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For a finite field \(\mathbb {F}_q\) and \(n \in \mathbb {N}\), the symplectic space is the vector space \(V = \mathbb {F}_q^n \times \mathbb {F}_q^n\), equipped with the symplectic form
A subspace \(S \le V\) is isotropic if \(\omega (u,v)=0\) for all \(u,v\in S\).
Classical codes \(C_Z, C_X \le \mathbb {F}_q^n\) satisfy the CSS condition if
Given \(C_Z \perp C_X\), the CSS stabilizer is the subspace of \(V\) generated by the \(X\)-type vectors \((x,0)\) for \(x \in C_Z\) and \(Z\)-type vectors \((0,z)\) for \(z \in C_X\).
The CSS stabilizer is isotropic if and only if the CSS condition holds.
If \(C_Z \perp C_X\), then \(\dim (\mathrm{cssStabilizer}) = \dim (C_Z) + \dim (C_X)\).
The CSS code \(Q(C_Z, C_X)\) encodes \(k = n - \dim (C_Z) - \dim (C_X)\) logical qudits, and has distance \(d = \min (d_Z, d_X)\) where \(d_Z\) (resp. \(d_X\)) is the minimum weight of a vector in \(C_Z^\perp \setminus C_X\) (resp. \(C_X^\perp \setminus C_Z\)).
For a real function \(f\) on the uniform Boolean cube \(\{ 0,1\} ^n\), write \(D_i f(x)=(f(x^{i\leftarrow 0})-f(x^{i\leftarrow 1}))/2\) and \(B(f)=\mathbb E\sqrt{\overline I(f)^2+\sum _i(D_i f)^2}\), where \(\overline I\) is the Gaussian profile extended by zero outside \((0,1)\). The Bobkov functional assigns the value \(B(f)\) to \(f\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. For real \(a,b\) with \(a+b,a-b\in [0,1]\), \(\overline I(a)\le \bigl(\sqrt{\overline I(a+b)^2+b^2}+\sqrt{\overline I(a-b)^2+b^2}\bigr)/2\).
For a real function \(f\) on the uniform Boolean cube \(\{ 0,1\} ^n\), write \(D_i f(x)=(f(x^{i\leftarrow 0})-f(x^{i\leftarrow 1}))/2\) and \(B(f)=\mathbb E\sqrt{\overline I(f)^2+\sum _i(D_i f)^2}\), where \(\overline I\) is the Gaussian profile extended by zero outside \((0,1)\). In dimension \(n=0\), \(B(f)=\overline I(\mathbb E f)\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(f:\{ 0,1\} ^{n+1}\to [0,1]\), put \(f_b(x)=f(x,b)\) for \(b\in \{ 0,1\} \), and set \(d=(f_0-f_1)/2\). Then \(\overline I(\mathbb E f)\le \bigl(\sqrt{\overline I(\mathbb E f_0)^2+(\mathbb E d)^2}+\sqrt{\overline I(\mathbb E f_1)^2+(\mathbb E d)^2}\bigr)/2\).
For a real function \(f\) on the uniform Boolean cube \(\{ 0,1\} ^n\), write \(D_i f(x)=(f(x^{i\leftarrow 0})-f(x^{i\leftarrow 1}))/2\) and \(B(f)=\mathbb E\sqrt{\overline I(f)^2+\sum _i(D_i f)^2}\), where \(\overline I\) is the Gaussian profile extended by zero outside \((0,1)\). If \(d\in \mathbb R\) and \(\overline I(\mathbb E f)\le B(f)\), then \(\sqrt{\overline I(\mathbb E f)^2+d^2}\le \sqrt{B(f)^2+d^2}\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(f:\{ 0,1\} ^{n+1}\to \mathbb R\), \(f_b(x)=f(x,b)\), and \(d=(f_0-f_1)/2\). For every \(b\in \{ 0,1\} \) and \(x\in \{ 0,1\} ^n\), \(\sqrt{(\sqrt{\overline I(f_b(x))^2+\sum _{i\lt n}(D_i f_b(x))^2})^2+d(x)^2}=\sqrt{\overline I(f(x,b))^2+\sum _{i\lt n+1}(D_i f(x,b))^2}\), where \(D_i\) denotes the coordinate half-difference.
For a real function \(h\) on a uniform Boolean cube, put \(B(h)=\mathbb E\sqrt{\overline I(h)^2+\sum _i(D_i h)^2}\), where \(D_i\) is the coordinate half-difference and \(\overline I\) is the Gaussian profile extended by zero outside \((0,1)\). For \(f:\{ 0,1\} ^{n+1}\to \mathbb R\), put \(f_b(x)=f(x,b)\) and \(d=(f_0-f_1)/2\). If \(\overline I(\mathbb E f_b)\le B(f_b)\), then \(\sqrt{\overline I(\mathbb E f_b)^2+(\mathbb E d)^2}\le \mathbb E_x\sqrt{\overline I(f(x,b))^2+\sum _{i\lt n+1}(D_i f(x,b))^2}\).
For a real function \(f\) on the uniform Boolean cube \(\{ 0,1\} ^n\), write \(D_i f(x)=(f(x^{i\leftarrow 0})-f(x^{i\leftarrow 1}))/2\) and \(B(f)=\mathbb E\sqrt{\overline I(f)^2+\sum _i(D_i f)^2}\), where \(\overline I\) is the Gaussian profile extended by zero outside \((0,1)\). For every \(n\ge 0\) and \(f:\{ 0,1\} ^n\to [0,1]\), \(\overline I(\mathbb E f)\le B(f)\).
For a real function \(f\) on the uniform Boolean cube \(\{ 0,1\} ^n\), write \(D_i f(x)=(f(x^{i\leftarrow 0})-f(x^{i\leftarrow 1}))/2\) and \(B(f)=\mathbb E\sqrt{\overline I(f)^2+\sum _i(D_i f)^2}\), where \(\overline I\) is the Gaussian profile extended by zero outside \((0,1)\). For \(f:\{ 0,1\} ^n\to [0,1]\) and \(0\le \rho \le 1\), \(B(T_\rho f)\le B(f)\), where \(T_\rho \) is the Boolean noise operator with correlation \(\rho \).
For a real function \(f\) on the uniform Boolean cube \(\{ 0,1\} ^n\), write \(D_i f(x)=(f(x^{i\leftarrow 0})-f(x^{i\leftarrow 1}))/2\) and \(B(f)=\mathbb E\sqrt{\overline I(f)^2+\sum _i(D_i f)^2}\), where \(\overline I\) is the Gaussian profile extended by zero outside \((0,1)\). For \(f:\{ 0,1\} ^n\to [0,1]\), the map \(\rho \mapsto B(T_\rho f)\) is increasing on \([0,1]\), where \(T_\rho \) is Boolean noise with correlation \(\rho \).
Let \(E\) be a real normed vector space and \(f:E\to \mathbb R\) be differentiable everywhere with derivative \(D(x):E\to \mathbb R\). Suppose \(x\mapsto D(x)\) is \(C\)-Lipschitz in operator norm, where \(C\ge 0\). For all \(x,y\in E\), \(|f(y)-f(x)-D(x)(y-x)|\le C\| y-x\| ^2\).
Let \(E\) be a real normed vector space and \(f:E\to \mathbb R\) be differentiable everywhere with derivative \(D(x):E\to \mathbb R\). Suppose \(x\mapsto D(x)\) is \(C\)-Lipschitz in operator norm, where \(C\ge 0\). For \(x,v\in E\) and \(h\in \mathbb R\), \(|(f(x+hv)-f(x))/2-(h/2)D(x)v|\le (C/2)h^2\| v\| ^2\).
For \(n,m\ge 0\) and \(x\in \{ 0,1\} ^{n\times m}\), define \(S_m(x)_k=m^{-1/2}\sum _{l=1}^m(-1)^{x_{kl}}\), interpreting division by zero as zero. If \(x^{ij\leftarrow b}\) replaces bit \((i,j)\) by \(b\), then \(S_m(x^{ij\leftarrow b})=S_m(x)+\bigl((-1)^b-(-1)^{x_{ij}}\bigr)m^{-1/2}e_i\), where \(e_i\) is the \(i\)th coordinate vector.
For \(n,m\ge 0\) and \(x\in \{ 0,1\} ^{n\times m}\), define \(S_m(x)_k=m^{-1/2}\sum _{l=1}^m(-1)^{x_{kl}}\), interpreting division by zero as zero. Let \(f:\mathbb R^n\to \mathbb R\) have derivative family \(D(z)\) that is globally \(C\)-Lipschitz in operator norm for the maximum norm on \(\mathbb R^n\), with \(C\ge 0\). For any bit \((i,j)\), \(\left|\frac{f(S_m(x^{ij\leftarrow 0}))-f(S_m(x^{ij\leftarrow 1}))}{2}-\frac{D(S_m(x))e_i}{\sqrt m}\right|\le 2C/m\).
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). The interpolation deficit is \(\Delta (t,c)=L(t,c)-L(1,c+b^2(1-t^2))\).
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). For every \(P,a,b,u,v\), the function \(c\mapsto L(1,c)\) is concave on \([0,\infty )\).
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). If \(P\) is continuous on \([0,1]\) and \(a+b,a-b\in [0,1]\), then \(t\mapsto L(t,0)\) is continuous on \([0,1]\).
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). For any real \(t\) and \(c\ge 0\), \(L(t,c)\le \sqrt c+L(t,0)\).
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). Assume \(P\ge 0\) on \([0,1]\) and, whenever \(z\pm d\in [0,1]\), \(P(z)\le \bigl(\sqrt{P(z+d)^2+d^2}+\sqrt{P(z-d)^2+d^2}\bigr)/2\). If \(a\pm b\in [0,1]\), \(0\le t\le 1\), and \(c\ge 0\), then \(L(t,c)\le \bigl(L(1,c+((1-t)b)^2)+L(2t-1,c+((1-t)b)^2)\bigr)/2\).
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). For every real \(t,c\), \(\sqrt c\le L(t,c)\), with the real square root taken to be zero on negative arguments.
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). If \(P\) is continuous on \([0,1]\) and \(a\pm b\in [0,1]\), the set \(\{ L(t,c)-L(1,c+b^2(1-t^2)):0\le t\le 1,\ c\ge 0\} \) is bounded above.
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). Assume \(P\ge 0\) on \([0,1]\) and, whenever \(z\pm d\in [0,1]\), \(P(z)\le \bigl(\sqrt{P(z+d)^2+d^2}+\sqrt{P(z-d)^2+d^2}\bigr)/2\). Put \(\Delta (t,c)=L(t,c)-L(1,c+b^2(1-t^2))\). If \(a\pm b\in [0,1]\), \(0\le t\le 1\), and \(c\ge 0\), then \(\Delta (t,c)\le \Delta (|2t-1|,c+((1-t)b)^2)/2\).
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). Assume \(P\ge 0\) on \([0,1]\) and, whenever \(z\pm d\in [0,1]\), \(P(z)\le \bigl(\sqrt{P(z+d)^2+d^2}+\sqrt{P(z-d)^2+d^2}\bigr)/2\). If \(P\) is also continuous on \([0,1]\), \(a\pm b\in [0,1]\), \(0\le t\le 1\), and \(c\ge 0\), then \(L(t,c)\le L(1,c+b^2(1-t^2))\).
For a profile \(P:\mathbb R\to \mathbb R\), \(a,b,t,c\in \mathbb R\), and \(u,v\in \mathbb R^m\), define \(L(t,c)=\frac12\bigl(\sqrt{P(a+tb)^2+c+\| u+tv\| _2^2}+\sqrt{P(a-tb)^2+c+\| u-tv\| _2^2}\bigr)\). Assume \(P\ge 0\) on \([0,1]\) and, whenever \(z\pm d\in [0,1]\), \(P(z)\le \bigl(\sqrt{P(z+d)^2+d^2}+\sqrt{P(z-d)^2+d^2}\bigr)/2\). If \(P\) is also continuous on \([0,1]\), \(a\pm b\in [0,1]\), and \(0\le t\le 1\), then \(L(t,(tb)^2)\le L(1,b^2)\).
Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). If \(f:\{ 0,1\} ^n\to [0,1]\) and \(0\le \rho \le 1\), then \((T_\rho f)(x)\in [0,1]\) for every \(x\).
Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For \(f:\{ 0,1\} ^{n+1}\to \mathbb R\), put \(f_b(x)=f(x,b)\), \(a_f=(f_0+f_1)/2\), and \(d_f=(f_0-f_1)/2\), and let \(s_0=1\), \(s_1=-1\). For every \(x,b\) and real \(\rho \), \((T_\rho f_b)(x)=(T_\rho a_f)(x)+s_b(T_\rho d_f)(x)\), where noise on the right acts on the first \(n\) coordinates.
Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For \(f:\{ 0,1\} ^{n+1}\to \mathbb R\), put \(f_b(x)=f(x,b)\), \(a_f=(f_0+f_1)/2\), and \(d_f=(f_0-f_1)/2\), and let \(s_0=1\), \(s_1=-1\). For every \(x,b\), \(D_n(T_\rho f)(x,b)=\rho (T_\rho d_f)(x)\), where \(D_n\) is the last-coordinate half-difference.
Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For \(f:\{ 0,1\} ^{n+1}\to \mathbb R\), put \(f_b(x)=f(x,b)\), \(a_f=(f_0+f_1)/2\), and \(d_f=(f_0-f_1)/2\), and let \(s_0=1\), \(s_1=-1\). For either bit \(b\), \((T_\rho f)_b=T_\rho a_f+s_b\rho T_\rho d_f\) as functions on \(\{ 0,1\} ^n\).
Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For \(f:\{ 0,1\} ^{n+1}\to \mathbb R\), put \(f_b(x)=f(x,b)\), \(a_f=(f_0+f_1)/2\), and \(d_f=(f_0-f_1)/2\), and let \(s_0=1\), \(s_1=-1\). For either bit \(b\), \(T_\rho f_b=T_\rho a_f+s_bT_\rho d_f\) as functions on \(\{ 0,1\} ^n\).
Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For \(f:\{ 0,1\} ^{n+1}\to \mathbb R\), put \(f_b(x)=f(x,b)\), \(a_f=(f_0+f_1)/2\), and \(d_f=(f_0-f_1)/2\), and let \(s_0=1\), \(s_1=-1\). If \(f\) takes values in \([0,1]\) and \(0\le \rho \le 1\), then \((T_\rho a_f)(x)\pm (T_\rho d_f)(x)\in [0,1]\) for every \(x\).
For \(P:\mathbb R\to \mathbb R\), \(f:\{ 0,1\} ^n\to \mathbb R\), and \(g=(g_1,\ldots ,g_m)\) of real functions on the cube, write \(V(f,g)(x)=(D_1f(x),\ldots ,D_nf(x),g_1(x),\ldots ,g_m(x))\) and \(F_P(f,g)=\mathbb E\sqrt{P(f)^2+\| V(f,g)\| _2^2}\), with \(D_i\) the coordinate half-difference. The extended functional is \(F_P(f,g)\) for arbitrary \(n,m\ge 0\).
For \(P:\mathbb R\to \mathbb R\), \(f:\{ 0,1\} ^n\to \mathbb R\), and \(g=(g_1,\ldots ,g_m)\) of real functions on the cube, write \(V(f,g)(x)=(D_1f(x),\ldots ,D_nf(x),g_1(x),\ldots ,g_m(x))\) and \(F_P(f,g)=\mathbb E\sqrt{P(f)^2+\| V(f,g)\| _2^2}\), with \(D_i\) the coordinate half-difference. For every \(x\), \(\| V(f,g)(x)\| _2^2=\sum _i(D_i f(x))^2+\sum _jg_j(x)^2\).
For a profile \(P:\mathbb R\to \mathbb R\), a real function \(h\) on a uniform Boolean cube, and a vector \(k\) of real functions on that cube, put \(V(h,k)(x)=((D_i h(x))_i,(k_j(x))_j)\) and \(F_P(h,k)=\mathbb E\sqrt{P(h)^2+\| V(h,k)\| _2^2}\), where \(D_i\) is the coordinate half-difference. Let \(f\) be a real function and \(g\) a vector of \(m\) real functions on \(\{ 0,1\} ^{n+1}\). Put \(f_b(x)=f(x,b)\), \(g_b(x)=g(x,b)\), \(a_f=(f_0+f_1)/2\), \(d_f=(f_0-f_1)/2\), and define \(a_g,d_g\) componentwise. For every \(x,b\), \(\sqrt{P(f(x,b))^2+\| V(f,g)(x,b)\| _2^2}=\sqrt{P(f_b(x))^2+d_f(x)^2+\| V(f_b,g_b)(x)\| _2^2}\).
For a profile \(P:\mathbb R\to \mathbb R\), a real function \(h\) on a uniform Boolean cube, and a vector \(k\) of real functions on that cube, put \(V(h,k)(x)=((D_i h(x))_i,(k_j(x))_j)\) and \(F_P(h,k)=\mathbb E\sqrt{P(h)^2+\| V(h,k)\| _2^2}\), where \(D_i\) is the coordinate half-difference. For real \(c,d\), functions \(f_1,f_2\) on the cube, and vectors \(g_1,g_2\) of \(m\) cube functions, \(V(cf_1+df_2,cg_1+dg_2)(x)=cV(f_1,g_1)(x)+dV(f_2,g_2)(x)\) for every \(x\).
For a profile \(P:\mathbb R\to \mathbb R\), a real function \(h\) on a uniform Boolean cube, and a vector \(k\) of real functions on that cube, put \(V(h,k)(x)=((D_i h(x))_i,(k_j(x))_j)\) and \(F_P(h,k)=\mathbb E\sqrt{P(h)^2+\| V(h,k)\| _2^2}\), where \(D_i\) is the coordinate half-difference. Let \(f\) be a real function and \(g\) a vector of \(m\) real functions on \(\{ 0,1\} ^{n+1}\). Put \(f_b(x)=f(x,b)\), \(g_b(x)=g(x,b)\), \(a_f=(f_0+f_1)/2\), \(d_f=(f_0-f_1)/2\), and define \(a_g,d_g\) componentwise. Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For \(s_0=1\), \(s_1=-1\), \(V(T_\rho f_b,T_\rho g_b)(x)=V(T_\rho a_f,T_\rho a_g)(x)+s_bV(T_\rho d_f,T_\rho d_g)(x)\) for every \(x,b\).
For a profile \(P:\mathbb R\to \mathbb R\), a real function \(h\) on a uniform Boolean cube, and a vector \(k\) of real functions on that cube, put \(V(h,k)(x)=((D_i h(x))_i,(k_j(x))_j)\) and \(F_P(h,k)=\mathbb E\sqrt{P(h)^2+\| V(h,k)\| _2^2}\), where \(D_i\) is the coordinate half-difference. Let \(f\) be a real function and \(g\) a vector of \(m\) real functions on \(\{ 0,1\} ^{n+1}\). Put \(f_b(x)=f(x,b)\), \(g_b(x)=g(x,b)\), \(a_f=(f_0+f_1)/2\), \(d_f=(f_0-f_1)/2\), and define \(a_g,d_g\) componentwise. Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For \(s_0=1\), \(s_1=-1\), \(V((T_\rho f)_b,(T_\rho g)_b)(x)=V(T_\rho a_f,T_\rho a_g)(x)+s_b\rho V(T_\rho d_f,T_\rho d_g)(x)\) for every \(x,b\).
For a profile \(P:\mathbb R\to \mathbb R\), a real function \(h\) on a uniform Boolean cube, and a vector \(k\) of real functions on that cube, put \(V(h,k)(x)=((D_i h(x))_i,(k_j(x))_j)\) and \(F_P(h,k)=\mathbb E\sqrt{P(h)^2+\| V(h,k)\| _2^2}\), where \(D_i\) is the coordinate half-difference. Let \(f\) be a real function and \(g\) a vector of \(m\) real functions on \(\{ 0,1\} ^{n+1}\). Put \(f_b(x)=f(x,b)\), \(g_b(x)=g(x,b)\), \(a_f=(f_0+f_1)/2\), \(d_f=(f_0-f_1)/2\), and define \(a_g,d_g\) componentwise. Then \(F_P(f,g)=\bigl(F_P(f_0,(d_f,g_0))+F_P(f_1,(d_f,g_1))\bigr)/2\), where \((d_f,g_b)\) adjoins \(d_f\) to the additional coordinates.
For a profile \(P:\mathbb R\to \mathbb R\), a real function \(h\) on a uniform Boolean cube, and a vector \(k\) of real functions on that cube, put \(V(h,k)(x)=((D_i h(x))_i,(k_j(x))_j)\) and \(F_P(h,k)=\mathbb E\sqrt{P(h)^2+\| V(h,k)\| _2^2}\), where \(D_i\) is the coordinate half-difference. Let \(f\) be a real function and \(g\) a vector of \(m\) real functions on \(\{ 0,1\} ^{n+1}\). Put \(f_b(x)=f(x,b)\), \(g_b(x)=g(x,b)\), \(a_f=(f_0+f_1)/2\), \(d_f=(f_0-f_1)/2\), and define \(a_g,d_g\) componentwise. Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For each \(x\), put \(a=T_\rho a_f(x)\), \(b=T_\rho d_f(x)\), \(u=V(T_\rho a_f,T_\rho a_g)(x)\), and \(v=V(T_\rho d_f,T_\rho d_g)(x)\). Then \(F_P(T_\rho f,T_\rho g)=\mathbb E_x\frac12\sum _{\epsilon \in \{ -1,1\} }\sqrt{P(a+\epsilon \rho b)^2+(\rho b)^2+\| u+\epsilon \rho v\| _2^2}\).
For a profile \(P:\mathbb R\to \mathbb R\), a real function \(h\) on a uniform Boolean cube, and a vector \(k\) of real functions on that cube, put \(V(h,k)(x)=((D_i h(x))_i,(k_j(x))_j)\) and \(F_P(h,k)=\mathbb E\sqrt{P(h)^2+\| V(h,k)\| _2^2}\), where \(D_i\) is the coordinate half-difference. Let \(f\) be a real function and \(g\) a vector of \(m\) real functions on \(\{ 0,1\} ^{n+1}\). Put \(f_b(x)=f(x,b)\), \(g_b(x)=g(x,b)\), \(a_f=(f_0+f_1)/2\), \(d_f=(f_0-f_1)/2\), and define \(a_g,d_g\) componentwise. Let \(T_\rho \) denote Boolean noise with real correlation \(\rho \), whose Fourier multiplier on a set of size \(k\) is \(\rho ^k\). For each \(x\), put \(a=T_\rho a_f(x)\), \(b=T_\rho d_f(x)\), \(u=V(T_\rho a_f,T_\rho a_g)(x)\), and \(v=V(T_\rho d_f,T_\rho d_g)(x)\). Then \(\mathbb E_x\frac12\sum _{\epsilon \in \{ -1,1\} }\sqrt{P(a+\epsilon b)^2+b^2+\| u+\epsilon v\| _2^2}=\bigl(F_P(T_\rho f_0,(T_\rho d_f,T_\rho g_0))+F_P(T_\rho f_1,(T_\rho d_f,T_\rho g_1))\bigr)/2\).
For \(P:\mathbb R\to \mathbb R\), \(f:\{ 0,1\} ^n\to \mathbb R\), and \(g=(g_1,\ldots ,g_m)\) of real functions on the cube, write \(V(f,g)(x)=(D_1f(x),\ldots ,D_nf(x),g_1(x),\ldots ,g_m(x))\) and \(F_P(f,g)=\mathbb E\sqrt{P(f)^2+\| V(f,g)\| _2^2}\), with \(D_i\) the coordinate half-difference. Assume \(P\ge 0\) on \([0,1]\) and, whenever \(z\pm d\in [0,1]\), \(P(z)\le \bigl(\sqrt{P(z+d)^2+d^2}+\sqrt{P(z-d)^2+d^2}\bigr)/2\). If \(P\) is continuous on \([0,1]\), \(f\) takes values in \([0,1]\), and \(0\le \rho \le 1\), then \(F_P(T_\rho f,T_\rho g)\le F_P(f,g)\), where \(T_\rho \) is Boolean noise applied to each component.
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(a,t,k\in \mathbb R\) and \(a\pm t\in (0,1)\), then the derivative of \(x\mapsto I(a+x)^2+kI(a-x)^2\) at \(t\) is \(-2I(a+t)Q(a+t)+2kI(a-t)Q(a-t)\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(a,t,k\in \mathbb R\) and \(a\pm t\in (0,1)\), then the derivative at \(t\) of \(x\mapsto -2I(a+x)Q(a+x)+2kI(a-x)Q(a-x)\) is \(2(Q(a+t)^2+kQ(a-t)^2-1-k)\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(b\ge 0\), \(a-b\gt 0\), \(a+b\lt 1\), and \(a\le 1/2\), then \(x\mapsto I(a+x)^2-I(a-x)^2\) is concave on \([0,b]\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(b\ge 0\), \(a-b\gt 0\), \(a+b\lt 1\), and \(a\le 1/2\), then \((I(a+b)^2-I(a-b)^2)^2\le 16I(a)^2Q(a)^2b^2\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(a,b\in \mathbb R\) satisfy \(a-b\gt 0\) and \(a+b\lt 1\), then \(x\mapsto I(a+x)^2+I(a-x)^2+2(1-Q(a)^2)x^2\) is convex on \([-b,b]\). Here \(Q(p)=0\) and \(I(p)=\varphi (0)\) outside \((0,1)\), and \([-b,b]\) may be empty.
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(b\ge 0\), \(a-b\gt 0\), and \(a+b\lt 1\), then \(2Q(a)^2b^2\le I(a+b)^2+I(a-b)^2+2b^2-2I(a)^2\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(b\ge 0\), \(a-b\gt 0\), \(a+b\lt 1\), and \(a\le 1/2\), then \(I(a)\le \bigl(\sqrt{I(a+b)^2+b^2}+\sqrt{I(a-b)^2+b^2}\bigr)/2\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(b\ge 0\) and \(a-b,a+b\in (0,1)\), then \(I(a)\le \bigl(\sqrt{I(a+b)^2+b^2}+\sqrt{I(a-b)^2+b^2}\bigr)/2\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(a,b\in \mathbb R\) and \(a-b,a+b\in (0,1)\), then \(I(a)\le \bigl(\sqrt{I(a+b)^2+b^2}+\sqrt{I(a-b)^2+b^2}\bigr)/2\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. If \(a,a+b\in [0,1]\), then \(\lim _{\rho \to 1^-}\overline I(a+\rho b)=\overline I(a+b)\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. If \(a\in (0,1)\) and \(a+b,a-b\in [0,1]\), then \(\overline I(a)\le \bigl(\sqrt{\overline I(a+b)^2+b^2}+\sqrt{\overline I(a-b)^2+b^2}\bigr)/2\).
Let \(I,J,B\) be finite sets, let \(B\) carry a measurable structure with measurable singletons, and let \(\mu _{ij}\) be sigma-finite measures on \(B\). The bijection \(B^{I\times J}\to (B^J)^I\) sending \(x\) to the array \(i\mapsto (j\mapsto x(i,j))\) pushes \(\bigotimes _{(i,j)}\mu _{ij}\) to \(\bigotimes _i\bigotimes _j\mu _{ij}\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(I\) be finite and \(f:\{ 0,1\} ^I\to [0,1]\). Under the independent uniform bit law, \(\overline I(\mathbb E f)\le \mathbb E\sqrt{\overline I(f)^2+\sum _{i\in I}(D_i f)^2}\), where \(D_i f(x)=(f(x^{i\leftarrow 0})-f(x^{i\leftarrow 1}))/2\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(I,J\) be finite and \(f:\{ 0,1\} ^{I\times J}\to [0,1]\). Under independent uniform bits, \(\overline I(\mathbb E f)\le \mathbb E\sqrt{\overline I(f)^2+\sum _{i\in I}\sum _{j\in J}(D_{ij}f)^2}\), where \(D_{ij}f(x)=(f(x^{ij\leftarrow 0})-f(x^{ij\leftarrow 1}))/2\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). For any set \(A\subseteq \mathbb R^n\) and real \(t\), its rotation sensitivity is the real value of the product-measure mass of the event \(\mathbf1_A(G)\ne \mathbf1_A(\cos t\, G+\sin t\, G\prime )\); for measurable \(A\), this is the disagreement probability.
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). For any \(A\subseteq \mathbb R^n\), its Gaussian surface area is the extended nonnegative quantity \(P_\gamma (A)=\liminf _{t\to 0^+}[\sqrt{\pi /2}\, s_A(t)/t]_+\), where \([u]_+=\max (u,0)\).
For a real number \(\alpha \), the closed Gaussian profile is \(U(\alpha )=\varphi (\Phi ^{-1}(\alpha ))\) if \(0\lt \alpha \lt 1\), and \(U(\alpha )=0\) otherwise, where \(\varphi \) and \(\Phi \) are the standard Gaussian density and distribution function.
Let \(n\ge 0\), \(m\gt 0\), \(P\in \mathbb R\), \(d\in \mathbb R^{n\times m}\), \(g\in \mathbb R^n\), and \(e\ge 0\). If \(|d_{ij}-g_i/\sqrt m|\le e\) for every \(i,j\), then \(\sqrt{P^2+\sum _{i,j}d_{ij}^2}\le \sqrt{P^2+\sum _i g_i^2}+\sqrt{nm}\, e\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(f:\mathbb R^n\to [0,1]\) be continuous, and let \(D(z)\) be a continuous family of real linear functionals on \(\mathbb R^n\) with \(\| D(z)\| \le M\) for all \(z\), for some real \(M\). Then \(z\mapsto \sqrt{\overline I(f(z))^2+\sum _i(D(z)e_i)^2}\) is a bounded continuous function.
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(f:\mathbb R^n\to [0,1]\) be differentiable with derivative family \(D(z)\), globally \(C\)-Lipschitz in operator norm for the maximum norm, where \(C\ge 0\). Write \(|D(z)|_2=\sqrt{\sum _{i=1}^n(D(z)e_i)^2}\). For independent uniform bits \(X_{ij}\), put \(S_r(X)_i=r^{-1/2}\sum _{j=1}^r(-1)^{X_{ij}}\). For every \(m\gt 0\), \(\overline I(\mathbb E f(S_m(X)))\le \mathbb E\sqrt{\overline I(f(S_m(X)))^2+|D(S_m(X))|_2^2}+\sqrt{nm}\, 2C/m\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). Let \(f:\mathbb R^n\to [0,1]\) be differentiable with derivative family \(D(z)\), globally \(C\)-Lipschitz in operator norm for the maximum norm, where \(C\ge 0\). Write \(|D(z)|_2=\sqrt{\sum _{i=1}^n(D(z)e_i)^2}\). If also \(\| D(z)\| \le M\) for every \(z\), for some \(M\in \mathbb R\), then \(\overline I(\int f\, d\gamma _n)\le \int \sqrt{\overline I(f(z))^2+|D(z)|_2^2}\, d\gamma _n(z)\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). Let \(D(z)\) be a continuous family of real linear functionals on \(\mathbb R^n\) and let \(M\in \mathbb R\) satisfy \(\| D(z)\| \le M\) for all \(z\). Then \(z\mapsto \sqrt{\sum _i(D(z)e_i)^2}\) is \(\gamma _n\)-integrable.
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). Let \(f:\mathbb R^n\to [0,1]\) be differentiable with derivative family \(D(z)\), globally \(C\)-Lipschitz in operator norm for the maximum norm, where \(C\ge 0\). Write \(|D(z)|_2=\sqrt{\sum _{i=1}^n(D(z)e_i)^2}\). If also \(\| D(z)\| \le M\) for every \(z\), for some \(M\in \mathbb R\), then \(\overline I(\int f\, d\gamma _n)\le \int \overline I(f)\, d\gamma _n+\int |D(z)|_2\, d\gamma _n(z)\).
Let \(A\subseteq \mathbb R\) be measurable, \(c,s\in \mathbb R\), and \(s\ne 0\). For \(Z\sim N(0,1)\) and the Gaussian density \(p_{s^2}(c,z)\) with mean \(c\) and variance \(s^2\), \(\mathbb E1_A(c+sZ)=\int _{\mathbb R}p_{s^2}(c,z)1_A(z)\, dz\).
Let \(v\gt 0\) and \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\). For every \(M\ge 0\), there exists a Lebesgue-integrable \(b:\mathbb R\to \mathbb R\) such that \(|((z-m)/v)p_v(m,z)|\le b(z)\) for all \(z\in \mathbb R\) and all \(m\) with \(|m|\le M\).
Let \(v\gt 0\) and \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\). Let \(f:\mathbb R\to \mathbb R\) be measurable with \(|f(z)|\le 1\) for all \(z\). For every \(c\in \mathbb R\), the derivative at \(c\) of \(m\mapsto \int p_v(m,z)f(z)\, dz\) is \(\int ((z-c)/v)p_v(c,z)f(z)\, dz\).
For \(v\ge 0\), put \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\) when \(v\gt 0\), and \(p_0(c,z)=0\); take division by zero to be zero. For real \(c,z\), the derivative at \(c\) of \(m\mapsto ((z-m)/v)p_v(m,z)\) is \(((z-c)^2/v^2-1/v)p_v(c,z)\).
Let \(v\gt 0\) and \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\). For every \(M\ge 0\), there is a Lebesgue-integrable \(b:\mathbb R\to \mathbb R\) such that \(|((z-m)^2/v^2-1/v)p_v(m,z)|\le b(z)\) for all real \(z\) and all \(m\) with \(|m|\le M\).
Let \(v\gt 0\) and \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\). If \(f:\mathbb R\to \mathbb R\) is measurable and \(|f|\le 1\), then for each real \(c\) the derivative at \(c\) of \(m\mapsto \int ((z-m)/v)p_v(m,z)f(z)\, dz\) is \(\int ((z-c)^2/v^2-1/v)p_v(c,z)f(z)\, dz\).
Let \(v\gt 0\) and \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\). Let \(c\in \mathbb R\) and \(g,b:\mathbb R\to \mathbb R\). If \(b\) is integrable for \(N(c,v)\) and \(|g(z)|\le b(z)\) for every \(z\), then \(|\int p_v(c,z)g(z)\, dz|\le \int b(z)\, dN(c,v)(z)\), with a nonintegrable integral on the left assigned value zero.
Let \(v\gt 0\) and \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\). For any function \(f:\mathbb R\to \mathbb R\) with \(|f|\le 1\) and any real \(c\), \(|\int ((z-c)^2/v^2-1/v)p_v(c,z)f(z)\, dz|\le 2/v\), where nonintegrable integrals are assigned zero.
Let \(v\gt 0\) and \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\). For any function \(f:\mathbb R\to \mathbb R\) with \(|f|\le 1\) and any real \(c\), \(|\int ((z-c)/v)p_v(c,z)f(z)\, dz|\le 1+1/v\), where nonintegrable integrals are assigned zero.
Let \(v\gt 0\) and \(p_v(c,z)=(2\pi v)^{-1/2}\exp (-(z-c)^2/(2v))\). For measurable \(f:\mathbb R\to \mathbb R\) with \(|f|\le 1\), the function \(c\mapsto \int p_v(c,z)f(z)\, dz\) is continuously differentiable on \(\mathbb R\).
Let \(\Phi ,\varphi \) denote the standard Gaussian distribution function and density. For \(t,y\in \mathbb R\), define \(q_{t,y}(\theta )=\Phi ((t-\sin \theta \, y)/\cos \theta )-\Phi (t)\). For every \(t,y\), \(q_{t,y}'(0)=-\varphi (t)y\).
Let \(\Phi ,\varphi \) denote the standard Gaussian distribution function and density. For \(t,y\in \mathbb R\), define \(q_{t,y}(\theta )=\Phi ((t-\sin \theta \, y)/\cos \theta )-\Phi (t)\). For every \(t,y\), \(|q_{t,y}(\theta )|/\theta \to \varphi (t)|y|\) as \(\theta \to 0^+\).
Let \(\Phi ,\varphi \) denote the standard Gaussian distribution function and density. For \(t,y\in \mathbb R\), define \(q_{t,y}(\theta )=\Phi ((t-\sin \theta \, y)/\cos \theta )-\Phi (t)\). For all \(t,y\in \mathbb R\) and \(0\lt \theta \lt 1\), \(|q_{t,y}(\theta )|/\theta \le 2\varphi (0)(|t|+|y|)\).
Let \(\Phi ,\varphi \) denote the standard Gaussian distribution function and density. For \(t,y\in \mathbb R\), define \(q_{t,y}(\theta )=\Phi ((t-\sin \theta \, y)/\cos \theta )-\Phi (t)\). For every real \(t\), if \(Y\sim N(0,1)\), then \(\mathbb E_Y(|q_{t,Y}(\theta )|/\theta )\to \varphi (t)\mathbb E|Y|\) as \(\theta \to 0^+\).
Let \(\Phi ,\varphi \) denote the standard Gaussian distribution function and density. For \(t,y\in \mathbb R\), define \(q_{t,y}(\theta )=\Phi ((t-\sin \theta \, y)/\cos \theta )-\Phi (t)\). If \(\cos \theta \gt 0\) and \(X,Y\) are independent standard Gaussian variables, then \(\Pr (1_{(-\infty ,t]}(X)\ne 1_{(-\infty ,t]}(\cos \theta \, X+\sin \theta \, Y))=\mathbb E_Y|q_{t,Y}(\theta )|\).
Let \(X,Y\) be independent standard Gaussian variables and \(\varphi \) their density. For every real \(t\), \(\sqrt{\pi /2}\, \Pr (1_{(-\infty ,t]}(X)\ne 1_{(-\infty ,t]}(\cos \theta \, X+\sin \theta \, Y))/\theta \to \varphi (t)\) as \(\theta \to 0^+\).
For a measurable \(A\subseteq \mathbb R^n\) and independent standard Gaussian vectors \(X,Y\), put \(R_A(\theta )=\Pr (1_A(X)\ne 1_A(\cos \theta \, X+\sin \theta \, Y))\). If \(a\in \mathbb R^n\) satisfies \(\sum _i a_i^2=1\) and \(H=\{ z:\sum _i a_i z_i\le t\} \), then \(R_H(\theta )=\Pr (1_{(-\infty ,t]}(U)\ne 1_{(-\infty ,t]}(\cos \theta \, U+\sin \theta \, V))\) for every real \(t,\theta \), where \(U,V\) are independent standard Gaussian scalars.
For a measurable \(A\subseteq \mathbb R^n\) and independent standard Gaussian vectors \(X,Y\), put \(R_A(\theta )=\Pr (1_A(X)\ne 1_A(\cos \theta \, X+\sin \theta \, Y))\). If \(a\in \mathbb R^n\) satisfies \(\sum _i a_i^2=1\) and \(H=\{ z:\sum _i a_i z_i\le t\} \) for \(t\in \mathbb R\), then \(\sqrt{\pi /2}\, R_H(\theta )/\theta \to \varphi (t)\) as \(\theta \to 0^+\), where \(\varphi \) is the standard Gaussian density.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). If \(A\) is measurable and \(\sin u\ne 0\), then for all \(x\in \mathbb R^n\), \(u_u(x)=\int _{\mathbb R^n}\prod _i p_{\sin ^2u}(\cos u\, x_i,z_i)1_A(z)\, dz\), where \(p_v(c,z)\) is the \(N(c,v)\) density.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). If \(A\) is measurable and \(\sin u\ne 0\), \(u_u\) is differentiable at every \(x\in \mathbb R^n\), with \(D u_u(x)[h]=\cos u\int 1_A(z)\, \partial _c[\prod _i p_{\sin ^2u}(c_i,z_i)]_{c=\cos u\, x}[h]\, dz\), where \(p_v\) is the Gaussian density and \(\partial _c\) is differentiation in the mean vector.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). If \(A\) is measurable and \(\sin u\ne 0\), then \(\| D u_u(x)\| \le n(1+1/\sin ^2u)\) for every \(x\in \mathbb R^n\), with operator norm for the maximum norm.
Let \(E\) be a real normed vector space, \(D:E\to \mathcal L(E,\mathbb R)\) be \(C\)-Lipschitz in operator norm with \(C\ge 0\), and \(L:E\to E\) be continuous linear with \(\| L\| \le 1\). Then \(x\mapsto D(Lx)\circ L\) is \(C\)-Lipschitz.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). If \(A\) is measurable and \(\sin u\ne 0\), then \(x\mapsto D u_u(x)\) is globally \(2n^2/\sin ^2u\)-Lipschitz in operator norm for the maximum norm on \(\mathbb R^n\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). The averaged indicator is the function \(u_\theta :\mathbb R^n\to \mathbb R\); for a nonintegrable integrand the integral is assigned zero.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). If \(A\) is measurable, then \(u_\theta (x)=\gamma _n\{ y:\cos \theta \, x+\sin \theta \, y\in A\} \) for every real \(\theta \) and every \(x\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). For independent \(X,Y\sim \gamma _n\), put \(R_A(\theta )=\Pr (1_A(X)\ne 1_A(\cos \theta \, X+\sin \theta \, Y))\). If \(A\) is measurable, then \(\int |u_\theta (x)-1_A(x)|\, d\gamma _n(x)=R_A(\theta )\) for every real \(\theta \).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). If \(A\) is measurable, then \(0\le u_\theta (x)\le 1\) for every real \(\theta \) and every \(x\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). If \(A\) is measurable, then \(\int |u_\theta -1_A|\, d\gamma _n\to 0\) as \(\theta \to 0\) through real values.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). For measurable \(A\) and every real \(\theta \), the function \(u_\theta \) is measurable.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). For measurable \(A\) and every real \(\theta \), the function \(u_\theta \) is \(\gamma _n\)-integrable.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). For measurable \(A\) and every real \(\theta \), \(\int u_\theta \, d\gamma _n=\gamma _n(A)\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For any real \(u\), if \(X_1,X_2,Y_1,Y_2\) are independent with law \(\gamma _n\), then \((\cos u\, X_1+\sin u\, Y_1,\cos u\, X_2+\sin u\, Y_2)\) has law \(\gamma _n\otimes \gamma _n\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). Let \(A\) be measurable, \(u,v\in \mathbb R\), and \(x,x\prime \in \mathbb R^n\). For independent \(Y,Y\prime \sim \gamma _n\), \(|u_u(x)-u_u(\cos v\, x+\sin v\, x\prime )|\le \Pr (1_A(\cos u\, x+\sin u\, Y)\ne 1_A(\cos u(\cos v\, x+\sin v\, x\prime )+\sin u(\cos v\, Y+\sin v\, Y\prime )))\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For \(A\subseteq \mathbb R^n\) and \(\theta \in \mathbb R\), put \(u_\theta (x)=\int 1_A(\cos \theta \, x+\sin \theta \, y)\, d\gamma _n(y)\). For independent \(X,Y\sim \gamma _n\), put \(R_A(\theta )=\Pr (1_A(X)\ne 1_A(\cos \theta \, X+\sin \theta \, Y))\). If \(A\) is measurable, then for every \(u,v\in \mathbb R\), \(\mathbb E_{X,Y}|u_u(X)-u_u(\cos v\, X+\sin v\, Y)|\le R_A(v)\).
Let \(X\) be a real random variable on a probability space, measurable up to a null set, with \(\mathbb EX=0\) and \(X(\omega )^2=1\) for every \(\omega \). Then \(X\) has finite fourth absolute moment and satisfies \(\mathbb EX^2=1\), \(\mathbb EX^3=0\), and \(\mathbb EX^4\le 9\), so it satisfies the invariance moment conditions.
Let \(B\) be a uniform bit and put \(X=1\) when \(B\) is false and \(X=-1\) when \(B\) is true. Then \(X\) satisfies the invariance moment conditions, including finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\).
Let \(X\) be a real function on a measured space \((\Omega ,\mu )\). It satisfies the invariance moment conditions when \(\mu \) is a probability measure, \(X\) is measurable up to a null set with finite fourth absolute moment, \(\mathbb E X=\mathbb E X^3=0\), \(\mathbb E X^2=1\), and \(\mathbb E X^4\le 9\).
Let \(X\) and \(Y\) be real random variables on measured spaces \((\Omega ,\mu )\) and \((\Xi ,\nu )\), where \(\nu \) is a probability measure. If \(X\) satisfies the invariance moment conditions and \(X\) and \(Y\) have the same distribution, then \(Y\) is measurable up to a null set with finite fourth absolute moment and moments \(\mathbb EY=\mathbb EY^3=0\), \(\mathbb EY^2=1\), and \(\mathbb EY^4\le 9\); thus \(Y\) satisfies the same moment conditions.
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{i\in S}x_i\) be a real multilinear polynomial of degree at most \(k\); let \(I_i(F)=\sum _{S\ni i}a_S^2\). Let \(X=(X_i)\) and \(Y=(Y_i)\) be families of independent real inputs on probability spaces \((\Omega ,\mu )\) and \((\Xi ,\nu )\), respectively. Each input is measurable up to a null set, has finite fourth absolute moment, and has first, second, and third moments \(0,1,0\) and fourth moment at most \(9\). If \(\psi \colon \mathbb R\to \mathbb R\) is \(C^4\) and \(C\ge 0\) satisfies \(|\psi ^{(4)}(t)|\le C\) for every \(t\), then \(|\mathbb E_\mu \psi (F(X))-\mathbb E_\nu \psi (F(Y))|\le (C/12)9^k\sum _i I_i(F)^2\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{i\in S}x_i\) be a real multilinear polynomial of degree at most \(k\); let \(I_i(F)=\sum _{S\ni i}a_S^2\). Let \(X=(X_i)\) and \(Y=(Y_i)\) be families of independent real inputs on probability spaces \((\Omega ,\mu )\) and \((\Xi ,\nu )\), respectively. Each input is measurable up to a null set, has finite fourth absolute moment, and has first, second, and third moments \(0,1,0\) and fourth moment at most \(9\). Suppose \(\sum _{S\ne \varnothing }a_S^2\le 1\), \(\varepsilon \ge 0\), and \(I_i(F)\le \varepsilon \) for every \(i\). If \(\psi \in C^4(\mathbb R)\) and \(C\ge 0\) bounds \(|\psi ^{(4)}|\) everywhere, then \(|\mathbb E_\mu \psi (F(X))-\mathbb E_\nu \psi (F(Y))|\le (C/12)k9^k\varepsilon \).
There exists a universal real constant \(K\gt 0\) with the following property. Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{i\in S}x_i\) be a real multilinear polynomial of degree at most \(k\); let \(I_i(F)=\sum _{S\ni i}a_S^2\). Let \(X=(X_i)\) and \(Y=(Y_i)\) be families of independent real inputs on probability spaces \((\Omega ,\mu )\) and \((\Xi ,\nu )\), respectively. Each input is measurable up to a null set, has finite fourth absolute moment, and has first, second, and third moments \(0,1,0\) and fourth moment at most \(9\). If \(\sum _{S\ne \varnothing }a_S^2\le 1\), \(\varepsilon \ge 0\), \(I_i(F)\le \varepsilon \) for every \(i\), and \(\psi \colon \mathbb R\to \mathbb R\) is \(c\)-Lipschitz with \(c\ge 0\), then \(|\mathbb E_\mu \psi (F(X))-\mathbb E_\nu \psi (F(Y))|\le Kc2^k\varepsilon ^{1/4}\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). The coordinate-free part at \(i\) is the multilinear polynomial with coefficient \(0\) on sets containing \(i\) and coefficient \(a_S\) on every set \(S\) not containing \(i\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). The coordinate derivative polynomial has, at each set \(S\), the Fourier coefficient indexed by \(S\) of the Boolean-cube derivative of \(F\) restricted to sign inputs.
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). The coefficient indexed by \(S\) in its coordinate derivative is \(0\) if \(i\in S\), and \(a_{S\cup \{ i\} }\) if \(i\notin S\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). Define \(A_i(x)=\sum _{S:i\notin S}a_S\prod _{j\in S}x_j\) and \(D_i(x)=\sum _{S:i\notin S}a_{S\cup \{ i\} }\prod _{j\in S}x_j\). If \(F\) has degree at most \(k+1\), where \(k\ge 0\) is an integer, then \(D_i\) has degree at most \(k\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). On \(\{ -1,1\} ^n\), evaluation of its coordinate derivative polynomial equals the discrete derivative of the cube function \(F\), where the discrete derivative at \(i\) is \(\tfrac 12(F|_{x_i=1}-F|_{x_i=-1})\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). Define \(A_i(x)=\sum _{S:i\notin S}a_S\prod _{j\in S}x_j\) and \(D_i(x)=\sum _{S:i\notin S}a_{S\cup \{ i\} }\prod _{j\in S}x_j\). The sum of the squared coefficients of \(D_i\) is the influence of coordinate \(i\) on the uniform sign-cube function \(F\): \(\sum _{S:i\notin S}a_{S\cup \{ i\} }^2=I_i(F)=\sum _{S\ni i}a_S^2\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). Let \(X_j\colon \Omega \to \mathbb R\) be arbitrary functions. For each \(\omega \in \Omega \), evaluation of the coordinate-free part of \(F\) on \(X(\omega )\) equals \(\sum _{S\subseteq \{ 1,\ldots ,n\} \setminus \{ i\} }a_S\prod _{j\in S}X_j(\omega )\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). Let \(X_j\colon \Omega \to \mathbb R\) be arbitrary functions. For each \(\omega \in \Omega \), evaluation of the coordinate derivative on \(X(\omega )\) equals \(\sum _{S\subseteq \{ 1,\ldots ,n\} \setminus \{ i\} }a_{S\cup \{ i\} }\prod _{j\in S}X_j(\omega )\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). Define \(A_i(x)=\sum _{S:i\notin S}a_S\prod _{j\in S}x_j\) and \(D_i(x)=\sum _{S:i\notin S}a_{S\cup \{ i\} }\prod _{j\in S}x_j\). For any functions \(X_j\colon \Omega \to \mathbb R\) and \(\omega \in \Omega \), \(F(X(\omega ))=A_i(X(\omega ))+X_i(\omega )D_i(X(\omega ))\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). Define \(A_i(x)=\sum _{S:i\notin S}a_S\prod _{j\in S}x_j\) and \(D_i(x)=\sum _{S:i\notin S}a_{S\cup \{ i\} }\prod _{j\in S}x_j\). If \((X_j)\) is an independent family of real functions on a measured space \((\Omega ,\mu )\) and each \(X_j\) is measurable up to a null set, then \(X_i\) is independent of the joint pair \((A_i(X),D_i(X))\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{i\in S}x_i\) be a real multilinear polynomial of degree at most \(k\); let \(I_i(F)=\sum _{S\ni i}a_S^2\). Let \((X_i)\) be independent real inputs on a probability space, each measurable up to a null set with finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\). Then \(\mathbb E[F(X)^4]\le 9^k(\sum _S a_S^2)^2\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). Define \(A_i(x)=\sum _{S:i\notin S}a_S\prod _{j\in S}x_j\) and \(D_i(x)=\sum _{S:i\notin S}a_{S\cup \{ i\} }\prod _{j\in S}x_j\). Suppose \(F\) has degree at most \(k+1\) and \((X_j)\) are independent real inputs on a probability space, each measurable up to a null set with finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\). Then \(\mathbb E[D_i(X)^4]\le 9^k I_i(F)^2\), where \(I_i(F)=\sum _{S\ni i}a_S^2\).
Let \(A,D,W\) be real functions on a probability space \((\Omega ,\mu )\) such that \((A,D)\) and \(W\) are measurable up to null sets and \(W\) is independent of \((A,D)\). If \(\psi \colon \mathbb R\to \mathbb R\) is continuous and \(\psi (A+DW)\) is integrable, then \(q(a,d)=\mathbb E_\mu [\psi (a+dW)]\) is integrable under the law \(\lambda \) of \((A,D)\) and \(\mathbb E_\mu [\psi (A+DW)]=\int q(a,d)\, d\lambda (a,d)\).
Let \(A,D,U,V\) be real random variables on a probability space, with \(A,D\in L^4\). Suppose each of \(U,V\) is independent of the pair \((A,D)\) and is measurable up to a null set with finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\). If \(\psi \in C^4(\mathbb R)\), \(C\ge 0\) bounds \(|\psi ^{(4)}|\) everywhere, and both \(\psi (A+DU)\) and \(\psi (A+DV)\) are integrable, then \(|\mathbb E\psi (A+DU)-\mathbb E\psi (A+DV)|\le (3C/4)\mathbb E[D^4]\).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{j\in S}x_j\) be a real multilinear polynomial and fix \(i\in \{ 1,\ldots ,n\} \). Suppose \(F\) has degree at most \(k+1\), and \(X,Y\) are real input families on one probability space, each independent within itself, agreeing as functions at every coordinate other than \(i\). Every input is measurable up to a null set with finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\). If \(\psi \in C^4(\mathbb R)\) and \(C\ge 0\) bounds \(|\psi ^{(4)}|\) everywhere, then \(|\mathbb E\psi (F(X))-\mathbb E\psi (F(Y))|\le (C/12)9^{k+1} I_i(F)^2\), where \(I_i(F)=\sum _{S\ni i}a_S^2\).
Let \(Z_{L,i},Z_{R,i}\colon \Omega \to \mathbb R\) be two indexed families, \(0\le i\lt n\), and let \(r\ge 0\) be an integer. The hybrid family \(H^{(r)}\) has \(H^{(r)}_i=Z_{R,i}\) for \(i\lt r\) and \(H^{(r)}_i=Z_{L,i}\) for \(i\ge r\).
Let \(Z_{L,i},Z_{R,i}\colon \Omega \to \mathbb R\) be two families indexed by \(0\le i\lt n\), and define \(H^{(r)}_i=Z_{R,i}\) if \(i\lt r\) and \(Z_{L,i}\) otherwise. Then \(H^{(n)}_i=Z_{R,i}\) for every \(i\), including the empty family when \(n=0\).
Let \(Z_{L,j},Z_{R,j}\colon \Omega \to \mathbb R\) be two families indexed by \(0\le j\lt n\), and put \(H^{(r)}_j=Z_{R,j}\) if \(j\lt r\) and \(Z_{L,j}\) otherwise. If \(i,j\in \{ 0,\ldots ,n-1\} \) and \(j\ne i\), then \(H^{(i)}_j=H^{(i+1)}_j\) as functions on \(\Omega \).
Let \(F(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }a_S\prod _{i\in S}x_i\) be a real multilinear polynomial of degree at most \(k\); let \(I_i(F)=\sum _{S\ni i}a_S^2\). Let \((Z_{L,i},Z_{R,i})_{i=1}^n\) be a mutually independent family of \(2n\) real inputs on a probability space, each measurable up to a null set with finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\). If \(\psi \in C^4(\mathbb R)\) and \(C\ge 0\) bounds \(|\psi ^{(4)}|\) everywhere, then \(|\mathbb E\psi (F(Z_L))-\mathbb E\psi (F(Z_R))|\le (C/12)9^k\sum _i I_i(F)^2\).
Let \(I\) be a finite index set and let \((X_i)_{i\in I}\) and \((Y_i)_{i\in I}\) be real input families on probability spaces \((\Omega ,\mu )\) and \((\Xi ,\nu )\), respectively. If each family is independent and \(X_i\) and \(Y_i\) are measurable up to null sets with the same distribution for every \(i\), then the random vectors \((X_i)_{i\in I}\) and \((Y_i)_{i\in I}\) have the same distribution.
For real \(\eta ,x\) and any real function \(\psi \), the interval smoothing is \(S_\eta \psi (x)=\eta ^{-1}\int _0^\eta \psi (x+t)\, dt\), with value zero for \(\eta =0\). The integral is oriented when \(\eta \lt 0\) and is assigned zero when its integrand is not integrable.
For \(\eta \gt 0\) and a real function \(\psi \), write \(S_\eta \psi (x)=\eta ^{-1}\int _0^\eta \psi (x+t)\, dt\). If \(\psi \) is continuous, then \(S_\eta \psi \) is differentiable at every real \(x\) with derivative \((\psi (x+\eta )-\psi (x))/\eta \).
For \(\eta \gt 0\) and a real function \(\psi \), write \(S_\eta \psi (x)=\eta ^{-1}\int _0^\eta \psi (x+t)\, dt\). For each integer \(m\ge 0\), if \(\psi \) is \(C^m\) on \(\mathbb R\), then \(S_\eta \psi \) is \(C^{m+1}\) on \(\mathbb R\).
For \(\eta \gt 0\) and a real function \(\psi \), write \(S_\eta \psi (x)=\eta ^{-1}\int _0^\eta \psi (x+t)\, dt\). If \(m\ge 0\) is an integer and \(\psi \in C^m(\mathbb R)\), then \((S_\eta \psi )^{(m+1)}(x)=(\psi ^{(m)}(x+\eta )-\psi ^{(m)}(x))/\eta \) for every real \(x\).
For \(\eta \gt 0\) and a real function \(\psi \), write \(S_\eta \psi (x)=\eta ^{-1}\int _0^\eta \psi (x+t)\, dt\). If \(\psi \) is \(c\)-Lipschitz for \(c\ge 0\), then \(|(S_\eta \psi )\prime (x)|\le c\) for every real \(x\).
For \(\eta \gt 0\) and a real function \(\psi \), write \(S_\eta \psi (x)=\eta ^{-1}\int _0^\eta \psi (x+t)\, dt\). If \(\psi \colon \mathbb R\to \mathbb R\) is \(c\)-Lipschitz with \(c\ge 0\), then \(S_\eta \psi \) is also \(c\)-Lipschitz.
For \(\eta \gt 0\) and a real function \(\psi \), write \(S_\eta \psi (x)=\eta ^{-1}\int _0^\eta \psi (x+t)\, dt\). If \(\psi \) is \(c\)-Lipschitz with \(c\ge 0\) and \(m\ge 0\) is an integer, put \(\phi =S_\eta ^{m+1}\psi \), where the exponent denotes repeated application of \(S_\eta \). Then \(\phi \) is \(C^{m+1}\) and \(c\)-Lipschitz, and for every real \(x\), \(|\phi (x)-\psi (x)|\le (m+1)c\eta \) and \(|\phi ^{(m+1)}(x)|\le c(2/\eta )^m\).
For every \(\eta \gt 0\), \(c\ge 0\), and \(c\)-Lipschitz function \(\psi \colon \mathbb R\to \mathbb R\), there exists \(\phi \in C^4(\mathbb R)\) such that \(|\phi (x)-\psi (x)|\le 4c\eta \) and \(|\phi ^{(4)}(x)|\le 8c/\eta ^3\) for all real \(x\).
Let \(f,g\colon \Omega \to \mathbb R\) be functions on a probability space \((\Omega ,\mu )\), with \(f\) integrable and \(g\) measurable up to a null set. If a real constant \(D\) satisfies \(|f(\omega )-g(\omega )|\le D\) for every \(\omega \), then \(|\int f\, d\mu -\int g\, d\mu |\le D\).
Let \(U\in L^4(\mu )\) and \(V\in L^4(\nu )\) be real inputs on probability spaces, and let \(B\ge 0\). Suppose that for every \(\phi \in C^4(\mathbb R)\) and \(C\ge 0\) bounding \(|\phi ^{(4)}|\) everywhere, \(|\mathbb E_\mu \phi (U)-\mathbb E_\nu \phi (V)|\le CB\). For every \(c\ge 0\), \(c\)-Lipschitz function \(\psi \colon \mathbb R\to \mathbb R\), and \(\eta \gt 0\), \(|\mathbb E_\mu \psi (U)-\mathbb E_\nu \psi (V)|\le 8c\eta +(8c/\eta ^3)B\).
Let \(U\in L^4(\mu )\) and \(V\in L^4(\nu )\) be real inputs on probability spaces, and let \(B\ge 0\). Suppose that for every \(\phi \in C^4(\mathbb R)\) and \(C\ge 0\) bounding \(|\phi ^{(4)}|\) everywhere, \(|\mathbb E_\mu \phi (U)-\mathbb E_\nu \phi (V)|\le CB\). For every \(c\ge 0\) and \(c\)-Lipschitz function \(\psi \colon \mathbb R\to \mathbb R\), \(|\mathbb E_\mu \psi (U)-\mathbb E_\nu \psi (V)|\le 16cB^{1/4}\), including \(B=0\).
Let \(Z\in L^4(\mu )\) be real-valued on a probability space \((\Omega ,\mu )\), and let \(\psi \colon \mathbb R\to \mathbb R\) be \(C^4\). If some \(C\ge 0\) bounds \(|\psi ^{(4)}(t)|\) for every real \(t\), then \(\psi (Z)\) is integrable under \(\mu \).
Let \(X\) be a real input on a probability space, measurable up to a null set with finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\). For any real function \(\psi \) and real \(a,b\), \(\mathbb E[\sum _{j=0}^3\psi ^{(j)}(a)(bX)^j/j!]=\psi (a)+\tfrac 12\psi ^{(2)}(a)b^2\). Here iterated differentiation assigns zero at a point where the derivative does not exist, so the identity requires no smoothness of \(\psi \).
Let \(X\) be a real input on a probability space, measurable up to a null set with finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\). Let \(\psi \colon \mathbb R\to \mathbb R\) be \(C^4\) and let \(C\ge 0\) satisfy \(|\psi ^{(4)}(t)|\le C\) for all real \(t\). For all real \(a,b\), \(|\mathbb E\psi (a+bX)-\psi (a)-\tfrac 12\psi ^{(2)}(a)b^2|\le (3C/8)b^4\).
Let \(X\) and \(Y\) be real inputs on possibly different probability spaces, each measurable up to a null set with finite fourth absolute moment, first three moments \(0,1,0\), and fourth moment at most \(9\). Let \(\psi \colon \mathbb R\to \mathbb R\) be \(C^4\) and let \(C\ge 0\) satisfy \(|\psi ^{(4)}(t)|\le C\) for all real \(t\). For every real \(a,b\), \(|\mathbb E\psi (a+bX)-\mathbb E\psi (a+bY)|\le (3C/4)b^4\).
Let \(\mu _m\) and \(\nu \) be probability measures on \(\mathbb R\). Assume that for every \(c\ge 0\) and every \(c\)-Lipschitz \(\psi :\mathbb R\to \mathbb R\), \(\int \psi \, d\mu _m\to \int \psi \, d\nu \). Then for each \(n\ge 0\) and every bounded continuous \(F:\mathbb R^n\to \mathbb R\), \(\int F\, d\mu _m^{\otimes n}\to \int F\, d\nu ^{\otimes n}\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). Let \(\nu _m\) be the law of the normalized sum of \(m+1\) independent uniform signs. For every \(n\ge 0\) and bounded continuous \(F:\mathbb R^n\to \mathbb R\), \(\int F\, d\nu _m^{\otimes n}\to \int F\, d\gamma _n\) as \(m\to \infty \).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For independent uniform bits \(X_{ij}\), put \(S_r(X)_i=r^{-1/2}\sum _{j=1}^r(-1)^{X_{ij}}\). For every \(n\ge 0\) and bounded continuous \(F:\mathbb R^n\to \mathbb R\), \(\mathbb E F(S_{m+1}(X))\to \int F\, d\gamma _n\) as \(m\to \infty \).
For \(n\ge 0\), define a multilinear coefficient family \(P_n\) on subsets of \(\{ 1,\ldots ,n\} \) by \(P_n(S)=1/\sqrt n\) when \(|S|=1\) and \(P_n(S)=0\) otherwise; division by zero is zero. Its evaluation is \(n^{-1/2}\sum _i z_i\).
Let \(n\ge 0\), let \(\Omega \) be a measurable space, and let \(X_i:\Omega \to \mathbb R\) be arbitrary functions. The multilinear polynomial with singleton coefficients \(1/\sqrt n\) and all other coefficients zero evaluates at \((X_i(\omega ))_i\) to \(\sum _iX_i(\omega )/\sqrt n\) for every \(\omega \in \Omega \), with division by zero equal to zero.
There is a universal \(K\gt 0\) such that, for every \(n\gt 0\), probability space \((\Omega ,\mu )\), and independent real random variables \(X_1,\ldots ,X_n\) with finite fourth moments, \(\mathbb E X_i=\mathbb E X_i^3=0\), \(\mathbb E X_i^2=1\), and \(\mathbb E X_i^4\le 9\), every \(c\)-Lipschitz \(\psi :\mathbb R\to \mathbb R\) with \(c\ge 0\) satisfies \(|\mathbb E\psi (n^{-1/2}\sum _iX_i)-\mathbb E\psi (Z)|\le Kc\, n^{-1/4}\), where \(Z\) is standard Gaussian.
There is a universal \(K\gt 0\) such that, for \(n\gt 0\), independent uniform signs \(\epsilon _1,\ldots ,\epsilon _n\), a standard Gaussian \(Z\), and any \(c\)-Lipschitz \(\psi :\mathbb R\to \mathbb R\) with \(c\ge 0\), \(|\mathbb E\psi (n^{-1/2}\sum _i\epsilon _i)-\mathbb E\psi (Z)|\le Kc\, n^{-1/4}\).
For every \(c\ge 0\) and \(c\)-Lipschitz \(\psi :\mathbb R\to \mathbb R\), if \(\epsilon _1,\ldots ,\epsilon _{m+1}\) are independent uniform signs and \(Z\) is standard Gaussian, then \(\mathbb E\psi ((m+1)^{-1/2}\sum _i\epsilon _i)\to \mathbb E\psi (Z)\) as \(m\to \infty \), where the first expectation is integration under the product sign law.
Let \(I\) be a finite set, \(c\in \mathbb R^I\), \(v\gt 0\), and \(\mu _{c,v}=\bigotimes _{i\in I}N(c_i,v)\). If \(f:\mathbb R^I\to \mathbb R\) is measurable with \(|f|\le 1\) and \(i\in I\), then \(z\mapsto ((z_i-c_i)/v)f(z)\) is \(\mu _{c,v}\)-integrable and \(|\int ((z_i-c_i)/v)f(z)\, d\mu _{c,v}(z)|\le 1+1/v\).
Let \(I\) be a set, \(c,z\in \mathbb R^I\), \(v\gt 0\), \(i,j\in I\), and \(a\in \mathbb R\) with \(|a|\le 1\). Then \(|(((z_i-c_i)(z_j-c_j)/v^2-\delta _{ij}/v)a)|\le ((z_i-c_i)^2+(z_j-c_j)^2)/(2v^2)+1/v\), where \(\delta _{ij}=1\) if \(i=j\) and zero otherwise.
Let \(I\) be a finite set, \(c\in \mathbb R^I\), \(v\gt 0\), and \(\mu _{c,v}=\bigotimes _{i\in I}N(c_i,v)\). For \(i,j\in I\), put \(H_{ij}(z)=(z_i-c_i)(z_j-c_j)/v^2-\delta _{ij}/v\), where \(\delta _{ij}\) is one if \(i=j\) and zero otherwise. If \(f:\mathbb R^I\to \mathbb R\) is measurable and \(|f|\le 1\), then \(H_{ij}f\) is \(\mu _{c,v}\)-integrable for every \(i,j\in I\).
Let \(I\) be a finite set, \(c\in \mathbb R^I\), \(v\gt 0\), and \(\mu _{c,v}=\bigotimes _{i\in I}N(c_i,v)\). For \(i,j\in I\), put \(H_{ij}(z)=(z_i-c_i)(z_j-c_j)/v^2-\delta _{ij}/v\), where \(\delta _{ij}\) is one if \(i=j\) and zero otherwise. For any \(f:\mathbb R^I\to \mathbb R\) with \(|f|\le 1\) and \(i,j\in I\), \(|\int H_{ij}f\, d\mu _{c,v}|\le 2/v\), assigning zero to a nonintegrable integral.
Let \(I\) be a finite set, \(c\in \mathbb R^I\), \(v\gt 0\), and \(\mu _{c,v}=\bigotimes _{i\in I}N(c_i,v)\). The measure \(\mu _{c,v}\) has Lebesgue density \(z\mapsto \prod _i p_v(c_i,z_i)\), where \(p_v(c_i,z_i)\) is the scalar Gaussian density.
Let \(I\) be a finite set, \(c\in \mathbb R^I\), \(v\gt 0\), and \(\mu _{c,v}=\bigotimes _{i\in I}N(c_i,v)\). For a real normed vector space \(E\) and any \(f:\mathbb R^I\to E\), \(\int f\, d\mu _{c,v}=\int _{\mathbb R^I}(\prod _i p_v(c_i,z_i))f(z)\, dz\), where \(p_v\) is the scalar Gaussian density and nonintegrable integrals on either side are assigned zero.
Let \(I\) be finite, \(c\in \mathbb R^I\), \(s\ne 0\), and \(f:\mathbb R^I\to \mathbb R\) measurable. For a standard Gaussian vector \(Y\), \(\mathbb E f(c+sY)=\int _{\mathbb R^I}(\prod _i p_{s^2}(c_i,z_i))f(z)\, dz\), where \(p_{s^2}\) is the scalar Gaussian density and nonintegrable integrals are assigned zero.
For \(v\gt 0\) and \(M\ge 0\), let \(p_v(m,z)\) be the Gaussian density with mean \(m\) and variance \(v\). There exists a nonnegative Lebesgue-integrable \(b:\mathbb R\to \mathbb R\) such that \(p_v(m,z)\le b(z)\), \(|((z-m)/v)p_v(m,z)|\le b(z)\), and \(|((z-m)^2/v^2-1/v)p_v(m,z)|\le b(z)\) whenever \(|m|\le M\).
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). The product-density derivative is \(A_v(c,z)\) for mean and input vectors \(c,z\in \mathbb R^I\).
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). For \(c,z,h\in \mathbb R^I\), \(A_v(c,z)[h]=G_v(c,z)\sum _i((z_i-c_i)/v)h_i\).
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). For every \(c,z\in \mathbb R^I\), the derivative at \(c\) of \(m\mapsto G_v(m,z)\) is \(A_v(c,z)\).
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). If \(v\gt 0\) and \(M\ge 0\), there exists a nonnegative Lebesgue-integrable \(b:\mathbb R^I\to \mathbb R\) such that \(|G_v(m,z)|\le b(z)\) and \(\| A_v(m,z)\| \le b(z)\) for all \(z\) and all \(m\) with \(|m_i|\le M\) for every \(i\), with operator norm for the maximum norm.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). For each fixed \(v\ge 0\), \((c,z)\mapsto A_v(c,z)\) is continuous in operator norm.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). For fixed \(v\ge 0\) and \(c\in \mathbb R^I\), \(z\mapsto A_v(c,z)\) is continuous in operator norm.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\) and \(c\in \mathbb R^I\), then \(z\mapsto f(z)A_v(c,z)\) is Lebesgue-integrable as a linear-functional-valued function.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\), the function \(m\mapsto \int _{\mathbb R^I}G_v(m,z)f(z)\, dz\) is differentiable at every \(c\), with derivative \(\int _{\mathbb R^I}f(z)A_v(c,z)\, dz\).
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\), the linear-functional-valued function \(c\mapsto \int _{\mathbb R^I}f(z)A_v(c,z)\, dz\) is continuous.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\), the map \(m\mapsto \int _{\mathbb R^I}G_v(m,z)f(z)\, dz\) is continuously differentiable.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\), \(c,h\in \mathbb R^I\), and \(Z\sim \bigotimes _iN(c_i,v)\), then \((\int f(z)A_v(c,z)\, dz)[h]=\sum _i\mathbb E[((Z_i-c_i)/v)f(Z)]h_i\).
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\) and \(c\in \mathbb R^I\), then \(\| \int f(z)A_v(c,z)\, dz\| \le |I|(1+1/v)\), in operator norm for the maximum norm.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. For \(c,z,h\in \mathbb R^I\), put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\). The mean score is the linear functional \(S_v(c,z)\).
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. For \(c,z,h\in \mathbb R^I\), put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\). For every \(c,z\), the derivative in \(c\) of \(S_v(c,z)\) is the bilinear map \((h,k)\mapsto -\sum _i h_i k_i/v\).
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). The product-density Hessian is the bilinear form \(T_v(c,z)\).
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). For all \(c,z\in \mathbb R^I\), the derivative at \(c\) of \(m\mapsto A_v(m,z)\) is \(T_v(c,z)\).
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). For all \(c,z,h,k\in \mathbb R^I\), \(T_v(c,z)[h,k]=G_v(c,z)\sum _{i,j}((z_i-c_i)(z_j-c_j)/v^2-\delta _{ij}/v)h_jk_i\), where \(\delta _{ij}\) is the Kronecker delta.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Assume \(c,z,a\in \mathbb R^I\), \(a_i\ge 0\), \(p_v(c_i,z_i)\le a_i\), and \(|((z_i-c_i)^2/v^2-1/v)p_v(c_i,z_i)|\le a_i\) for every \(i\in I\). For every \(i\in I\), \(((z_i-c_i)^2/v^2)G_v(c,z)\le (1+1/v)\prod _j a_j\).
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Assume \(c,z,a\in \mathbb R^I\), \(a_i\ge 0\), \(p_v(c_i,z_i)\le a_i\), and \(|((z_i-c_i)^2/v^2-1/v)p_v(c_i,z_i)|\le a_i\) for every \(i\in I\). If \(v\gt 0\), then for every \(i,j\in I\), \(|((z_i-c_i)(z_j-c_j)/v^2-\delta _{ij}/v)G_v(c,z)|\le (1+2/v)\prod _k a_k\), where \(\delta _{ij}\) is the Kronecker delta.
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). For fixed \(v\ge 0\), the map \((c,z)\mapsto T_v(c,z)\) is continuous in bilinear operator norm.
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). If \(v\gt 0\) and \(M\ge 0\), there is a nonnegative Lebesgue-integrable \(b:\mathbb R^I\to \mathbb R\) such that \(\| T_v(m,z)\| \le b(z)\) for every \(z\) and every \(m\) satisfying \(|m_i|\le M\) for all \(i\); the bilinear operator norm uses maximum norms.
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\) and \(c\in \mathbb R^I\), then \(z\mapsto f(z)T_v(c,z)\) is Lebesgue-integrable as a bilinear-form-valued function.
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\), then \(m\mapsto \int f(z)A_v(m,z)\, dz\) is differentiable at every \(c\), with derivative \(\int f(z)T_v(c,z)\, dz\).
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\), \(c,h,k\in \mathbb R^I\), and \(Z\sim \bigotimes _iN(c_i,v)\), then \((\int f(z)T_v(c,z)\, dz)[h,k]=\sum _{i,j}\mathbb E[((Z_i-c_i)(Z_j-c_j)/v^2-\delta _{ij}/v)f(Z)]h_jk_i\), where \(\delta _{ij}\) is the Kronecker delta.
Let \(I\) be finite and \(v\ge 0\). Put \(G_v(c,z)=\prod _i p_v(c_i,z_i)\), where \(p_v\) is the Gaussian density of variance \(v\) for \(v\gt 0\) and \(p_0=0\), and put \(S_v(c,z)[h]=\sum _i((z_i-c_i)/v)h_i\), with division by zero equal to zero. Define \(A_v(c,z)[h]=G_v(c,z)S_v(c,z)[h]\) and \(T_v(c,z)[h,k]=G_v(c,z)(S_v(c,z)[h]S_v(c,z)[k]-\sum _i h_i k_i/v)\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\), then \(\| \int f(z)T_v(c,z)\, dz\| \le 2|I|^2/v\) for every \(c\in \mathbb R^I\), using maximum norms for the bilinear operator norm.
Let \(I\) be finite and \(v\ge 0\). Write \(p_v(c,z)\) for the \(N(c,v)\) density when \(v\gt 0\) and set \(p_0(c,z)=0\); define \(G_v(c,z)=\prod _{i\in I}p_v(c_i,z_i)\) and interpret division by zero as zero. Let \(A_v(c,z)\) be the linear functional \(h\mapsto \sum _i\bigl(\prod _{j\ne i}p_v(c_j,z_j)\bigr)((z_i-c_i)/v)p_v(c_i,z_i)h_i\). Let \(f:\mathbb R^I\to \mathbb R\) be measurable with \(|f|\le 1\). If \(v\gt 0\), the map \(c\mapsto \int f(z)A_v(c,z)\, dz\) is globally \(2|I|^2/v\)-Lipschitz, with maximum norm on mean vectors and operator norm on linear functionals.
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). For every \(a\in (0,1)\), the derivative of \(p\mapsto -Q(p)\) at \(a\) is \(-1/I(a)\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). If \(a,b\in (0,1)\), \(a\le b\), and \(a+b\le 1\), then \(Q(b)^2\le Q(a)^2\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). For \(a\in (0,1)\), the derivative of \(p\mapsto I(p)Q(p)\) at \(a\) is \(1-Q(a)^2\).
Let \(\Phi \) and \(\varphi \) be the standard Gaussian distribution function and density, and put \(Q=\Phi ^{-1}\) and \(I(p)=\varphi (Q(p))\) for \(0\lt p\lt 1\). For \(a\in (0,1)\), the derivative of \(p\mapsto I(p)^2\) at \(a\) is \(-2I(a)Q(a)\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. If \(a\in (0,1)\), then \(\overline I(a)=\varphi (\Phi ^{-1}(a))\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. For every real \(a\), \(\overline I(1-a)=\overline I(a)\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. For \(a\in (0,1)\), \(\overline I'(a)=-\Phi ^{-1}(a)\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. The function \(\overline I\) is continuous on \([0,1]\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. The function \(\overline I\) is continuous on \(\mathbb R\).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. If \(\mu \) is a finite measure on a measurable space \(\Omega \) and \(f:\Omega \to \mathbb R\) is measurable with \(f(x)\in [0,1]\) for every \(x\), then \(\overline I\circ f\) is \(\mu \)-integrable.
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. For every \(\varepsilon \gt 0\) there exists \(C\gt 0\) such that \(\overline I(p)\le \varepsilon +C|p-b|\) for all \(p\in [0,1]\) and \(b\in \{ 0,1\} \).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For a measurable \(A\subseteq \mathbb R^n\), let \(u_\theta (x)=\mathbb E_Y1_A(\cos \theta \, x+\sin \theta \, Y)\) with \(Y\sim \gamma _n\), and put \(R_A(\theta )=\Pr (1_A(X)\ne 1_A(\cos \theta \, X+\sin \theta \, Y))\) for independent \(X,Y\sim \gamma _n\). For every \(\varepsilon \gt 0\) there exists \(C\gt 0\) such that \(\int \overline I(u_\theta (x))\, d\gamma _n(x)\le \varepsilon +CR_A(\theta )\) for every real \(\theta \).
Let \(\Phi ,\varphi \) be the standard Gaussian distribution function and density, and let \(\overline I(p)=\varphi (\Phi ^{-1}(p))\) for \(0\lt p\lt 1\), with \(\overline I(p)=0\) elsewhere. Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\). For a measurable \(A\subseteq \mathbb R^n\), let \(u_\theta (x)=\mathbb E_Y1_A(\cos \theta \, x+\sin \theta \, Y)\) with \(Y\sim \gamma _n\), and put \(R_A(\theta )=\Pr (1_A(X)\ne 1_A(\cos \theta \, X+\sin \theta \, Y))\) for independent \(X,Y\sim \gamma _n\). Then \(\int \overline I(u_\theta (x))\, d\gamma _n(x)\to 0\) as \(\theta \to 0\) through real values.
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). For any real \(t\), the map \((g,h)\mapsto (\cos t\, g+\sin t\, h,-\sin t\, g+\cos t\, h)\) preserves \(\gamma _n\otimes \gamma _n\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). If \(A\subseteq \mathbb R^n\) is measurable, then \(s_A(s+t)\le s_A(s)+s_A(t)\) for all real \(s,t\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). For any measurable \(A\subseteq \mathbb R^n\), integer \(\ell \ge 0\), and real angles \(\delta _1,\ldots ,\delta _\ell \), \(s_A(\sum _{j=1}^\ell \delta _j)\le \sum _{j=1}^\ell s_A(\delta _j)\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). For any measurable \(A\subseteq \mathbb R^n\), real \(\delta \), and positive integer \(\ell \), \(s_A(\delta )\le \ell \, s_A(\delta /\ell )\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). If \(A\subseteq \mathbb R^n\) is measurable and \(\alpha =\gamma _n(A)\), then \(s_A(\pi /2)=2\alpha (1-\alpha )\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). If \(A\subseteq \mathbb R^n\) is measurable with \(\gamma _n(A)=1/2\), then \(s_A(\pi /(2\ell ))\ge 1/(2\ell )\) for every positive integer \(\ell \).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). For any measurable \(A\subseteq \mathbb R^n\), \(s_A(t)\to 0\) as the real angle \(t\to 0\).
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\) and \(X,Y\) independent with that law. For a linear functional \(L:\mathbb R^n\to \mathbb R\), put \(|L|_2=\sqrt{\sum _i L(e_i)^2}\), where \(e_i\) is the canonical basis. Then \(\mathbb E|L(X)|=|L|_2\, 2\varphi (0)\), where \(\varphi \) is the one-dimensional standard Gaussian density.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\) and \(X,Y\) independent with that law. Let \(f:\mathbb R^n\to \mathbb R\) be differentiable everywhere with derivative \(L(x)\), and put \(|L(x)|_2=\sqrt{\sum _i(L(x)e_i)^2}\), where \(e_i\) is the canonical basis. Let \(C\in [0,\infty ]\). If \(\mathbb E(|f(\cos \theta \, X+\sin \theta \, Y)-f(X)|/\theta )\le C\) for every \(\theta \gt 0\), then \(\mathbb E|L(X)Y|\le C\), with both expectations interpreted as nonnegative extended integrals.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\) and \(X,Y\) independent with that law. For a linear functional \(L:\mathbb R^n\to \mathbb R\), put \(|L|_2=\sqrt{\sum _i L(e_i)^2}\), where \(e_i\) is the canonical basis. The nonnegative extended expectation satisfies \(\mathbb E|L(X)|=|L|_2\, 2\varphi (0)\), with the finite expression on the right embedded into \([0,\infty ]\) and \(\varphi \) the standard Gaussian density.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\) and \(X,Y\) independent with that law. If \(L:\mathbb R^n\to \mathcal L(\mathbb R^n,\mathbb R)\) is continuous, put \(|L(x)|_2=\sqrt{\sum _i(L(x)e_i)^2}\). Then \(\mathbb E|L(X)Y|=(\mathbb E|L(X)|_2)\, 2\varphi (0)\) as nonnegative extended integrals, where \(\varphi \) is the standard Gaussian density.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\) and \(X,Y\) independent with that law. Let \(f:\mathbb R^n\to \mathbb R\) be differentiable everywhere with derivative \(L(x)\), and put \(|L(x)|_2=\sqrt{\sum _i(L(x)e_i)^2}\), where \(e_i\) is the canonical basis. Assume \(L\) is continuous and \(C\in [0,\infty ]\). If \(\mathbb E(\sqrt{\pi /2}\, |f(\cos \theta \, X+\sin \theta \, Y)-f(X)|/\theta )\le C\) for every \(\theta \gt 0\), then \(\mathbb E|L(X)|_2\le C\), with all expectations nonnegative extended integrals.
Let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\) and \(X,Y\) independent with that law. Let \(f:\mathbb R^n\to \mathbb R\) be differentiable everywhere with derivative \(L(x)\), and put \(|L(x)|_2=\sqrt{\sum _i(L(x)e_i)^2}\), where \(e_i\) is the canonical basis. Assume \(L\) is continuous, \(0\le f\le 1\), and \(C\in [0,\infty ]\). If \(\sqrt{\pi /2}\, \mathbb E|f(\cos \theta \, X+\sin \theta \, Y)-f(X)|/\theta \le C\) for every \(\theta \gt 0\), where the real expectation is embedded into \([0,\infty ]\), then the nonnegative extended expectation satisfies \(\mathbb E|L(X)|_2\le C\).
Let \(f\colon \mathbb R\to \mathbb R\) be continuous at zero with \(f(0)=0\) and \(f(s+t)\le f(s)+f(t)\) for all real \(s,t\). Then for every \(\delta \gt 0\), \([f(\delta )/\delta ]_+\le \liminf _{t\to 0^+}[f(t)/t]_+\) in \([0,\infty ]\), where \([u]_+=\max (u,0)\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). If \(A\subseteq \mathbb R^n\) is measurable and \(\delta \gt 0\), then \([\sqrt{\pi /2}\, s_A(\delta )/\delta ]_+\le P_\gamma (A)\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). For every measurable \(A\subseteq \mathbb R^n\), \([\sqrt{\pi /2}\, s_A(t)/t]_+\to P_\gamma (A)\) as \(t\to 0^+\), with convergence in \([0,\infty ]\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). For any sets \(A,B\subseteq \mathbb R^n\) with \(\mathbf1_A=\mathbf1_B\) almost everywhere under \(\gamma _n\), \(s_A(t)=s_B(t)\) for every real \(t\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). If measurable sets \(A,B\subseteq \mathbb R^n\) agree almost everywhere under \(\gamma _n\), then \(P_\gamma (A)=P_\gamma (B)\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). For a vector \(v\in \mathbb R^n\) with \(\sum _i v_i^2=1\) and real \(t\), the halfspace \(H=\{ z:\sum _i v_i z_i\le t\} \) satisfies \(P_\gamma (H)=\varphi (t)\), where \(\varphi (t)=(2\pi )^{-1/2}e^{-t^2/2}\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). For every measurable \(A\subseteq \mathbb R^n\) and \(\delta \gt 0\), \([\sqrt{\pi /2}\, s_A(\delta )/\delta ]_+\le P_\gamma (A)\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). If \(A\subseteq \mathbb R^n\) is measurable with \(\alpha =\gamma _n(A)\), then \(2\sqrt{2/\pi }\, \alpha (1-\alpha )\le P_\gamma (A)\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). Let \(A\subseteq \mathbb R^n\) be measurable and \(f\colon \mathbb R^n\to [0,1]\) be differentiable with continuous derivative \(D(x)\). If \(\mathbb E|f(G)-f(\cos \theta \, G+\sin \theta \, G\prime )|\le s_A(\theta )\) for every real \(\theta \), then \(\int \| \nabla f(x)\| _2\, d\gamma _n(x)\le P_\gamma (A)\) as a nonnegative extended integral.
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). Let \(A\subseteq \mathbb R^n\) be measurable, \(0\lt \alpha \lt 1\), and let \(U\) be the standard Gaussian isoperimetric profile extended by zero outside \((0,1)\). Set \(f_u(x)=\mathbb E[\mathbf1_A(\cos u\, x+\sin u\, G)]\). If \(U(\alpha )\le [\int U(f_u)\, d\gamma _n]_++P_\gamma (A)\) for every \(0\lt u\lt \pi /2\), then \(U(\alpha )\le P_\gamma (A)\).
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). Let \(A\subseteq \mathbb R^n\) be measurable with \(\gamma _n(A)=\alpha \in (0,1)\) and put \(f_u(x)=\mathbb E[\mathbf1_A(\cos u\, x+\sin u\, G)]\). Suppose that for each \(0\lt u\lt \pi /2\), \(f_u\) has a derivative family \(D_u\) and there exist constants \(C_u\ge 0\) and \(M_u\in \mathbb R\) such that \(D_u\) is \(C_u\)-Lipschitz and \(\| D_u(x)\| \le M_u\) for all \(x\). Then \(\varphi (\Phi ^{-1}(\alpha ))\le P_\gamma (A)\), where \(\varphi \) and \(\Phi \) are the standard Gaussian density and distribution function.
Let \(n\ge 0\), let \(\gamma _n\) be standard Gaussian measure on \(\mathbb R^n\), and let \(G,G\prime \) be independent vectors with law \(\gamma _n\). Write \(s_A(t)\) for the real value of the product-measure mass \((\gamma _n\otimes \gamma _n)\{ (g,h):\mathbf1_A(g)\ne \mathbf1_A(\cos t\, g+\sin t\, h)\} \). Write \(P_\gamma (A)=\liminf _{t\to 0^+}\bigl[\sqrt{\pi /2}\, s_A(t)/t\bigr]_+\) in \([0,\infty ]\), where \([x]_+=\max (x,0)\). For every measurable \(A\subseteq \mathbb R^n\) of Gaussian volume \(\alpha \in (0,1)\), \(\varphi (\Phi ^{-1}(\alpha ))\le P_\gamma (A)\), where \(\varphi \) and \(\Phi \) are the standard Gaussian density and distribution function.
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube, \(i\in [n]\), and \(p\gt 0\). Then \(\mathbb E|D_if|^p=I_i(f)\), where \(D_if(x)=(f(x^{i\leftarrow 0})-f(x^{i\leftarrow 1}))/2\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube, \(i\in [n]\), and \(0\le \rho \le 1\). Then \(I_i^{(\rho )}(f)\le I_i(f)^{2/(1+\rho )}\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube and \(i\in [n]\). Then \(I_i^{(1/3)}(f)\le I_i(f)^{3/2}\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube with variance \(V\gt 0\), and put \(a=I(f)/V\). Then \((9/a^2)9^{-a}\le \max _{i\in [n]}|I_i(f)|\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube in dimension \(n\ge 2\). Some coordinate \(i\) satisfies \(\operatorname {Var}(f)\log n/(30n)\le I_i(f)\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
There exists an absolute \(c\gt 0\) such that for every natural dimension \(n\) and every \(f:\{ 0,1\} ^n\to \{ -1,1\} \) under the uniform law, \(c\operatorname {Var}(f)(\log n/n)\le \sup _{i\in [n]}|I_i(f)|\), with the empty supremum and division by zero interpreted as zero. For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
There exists an absolute \(C\gt 0\) such that for every \(f:\{ 0,1\} ^n\to \{ -1,1\} \) under the uniform law with \(\operatorname {Var}(f)\ge 1/2\), there is a nonempty \(S\subseteq [n]\) with \(|S|\le CI(f)\) and \(\widehat f(S)^2\ge \exp (-CI(f)^2)\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube, \(J\subseteq [n]\), \(k\in \mathbb N\), \(0\le \rho \le 1\), and \(\tau \ge 0\). If \(I_i(f)\le \tau \) for every \(i\notin J\), then \(\rho ^k\sum _{|S|\le k,\ S\nsubseteq J}\widehat f(S)^2\le \tau ^{(1-\rho )/(1+\rho )}I(f)\), with \(0^0=1\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(d\) be a DNF on \(n\) variables of width at most a positive integer \(w\). Suppose every term has no repeated literals and any two literals in a term sharing a variable are equal. For every \(t\in \mathbb N\), its sign encoding \(f(x)=(-1)^{d(x)}\) satisfies \(\sum _{|S|\gt 120w(t+1)}\widehat f(S)^2\le 8\, 2^{-t}\) under the uniform cube law.
Let \(d\) be a DNF on \(n\) variables of width at most a positive integer \(w\), and let \(t\in \mathbb N\). Its sign encoding \(f(x)=(-1)^{d(x)}\) satisfies \(\sum _{|S|\gt 120w(t+1)}\widehat f(S)^2\le 8\, 2^{-t}\) under the uniform cube law.
For every DNF \(d\) on \(n\) variables, the total influence under the uniform cube law of \(x\mapsto (-1)^{d(x)}\) is at most twice the width of \(d\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
There exists an absolute \(C\gt 0\) such that for every positive integer \(w\) and \(0\lt \varepsilon \le 1/2\), there exists \(t\in \mathbb N\) satisfying \(8\, 2^{-t}\le \varepsilon \) and \(((2w)^3/\varepsilon ^2)9^{120w(t+1)}\le (1/\varepsilon )^{Cw}\).
There exists an absolute \(C\gt 0\) such that for every DNF \(d\) on \(n\) variables of width at most \(w\in \mathbb N\) and every \(0\lt \varepsilon \le 1\), there exist \(J\subseteq [n]\) and \(h:\{ 0,1\} ^n\to \{ -1,1\} \) depending only on \(J\), with \(|J|\le (1/\varepsilon )^{Cw}\) and \(\Pr _x((-1)^{d(x)}\ne h(x))\le \varepsilon \) under the uniform cube law.
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube, \(0\lt \varepsilon \le 1\), and \(k\in \mathbb N\). Write \(I=I(f)\), put \(\tau =1\) if \(I=0\) and \(\tau =\varepsilon ^2/(I^2 9^k)\) otherwise, and set \(J=\{ i:I_i(f)\ge \tau \} \). Then \(|J|\le I^3 9^k/\varepsilon ^2\) and \(\sum _{|S|\le k,\ S\nsubseteq J}\widehat f(S)^2\le \varepsilon \); if \(\sum _{|S|\gt k}\widehat f(S)^2\le \varepsilon \), then \(\sum _{\neg (S\subseteq J\ \mathrm{and}\ |S|\le k)}\widehat f(S)^2\le 2\varepsilon \) and there exists \(h:\{ 0,1\} ^n\to \{ -1,1\} \) depending only on \(J\) with \(\Pr (f\ne h)\le \varepsilon \). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
There exists an absolute \(C\gt 0\) such that for every \(f:\{ 0,1\} ^n\to \{ -1,1\} \) under the uniform law and \(0\lt \varepsilon \le 1\), there exist \(J\subseteq [n]\) and \(h:\{ 0,1\} ^n\to \{ -1,1\} \) depending only on \(J\) with \(|J|\le \exp (CI(f)/\varepsilon )\), \(\sum _{\neg (S\subseteq J\ \mathrm{and}\ |S|\le I(f)/\varepsilon )}\widehat f(S)^2\le 2\varepsilon \), and \(\Pr (f\ne h)\le \varepsilon \). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube with \(I=I(f)\gt 0\). For \(\varepsilon ,\eta \gt 0\) and \(k\in \mathbb N\) satisfying \(\sum _{|S|\gt k}\widehat f(S)^2\le \varepsilon \), put \(\rho =\eta /(2+\eta )\). There exist \(J\subseteq [n]\) and \(h:\{ 0,1\} ^n\to \{ -1,1\} \) depending only on \(J\), with \(|J|\le I^{2+\eta }/(\varepsilon \rho ^k)^{1+\eta }\) and \(\Pr (f\ne h)\le \varepsilon \). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube, \(J\subseteq [n]\), and \(x\in \{ 0,1\} ^n\). If \(x\star _Jy\) keeps \(x\) on \(J\) and \(y\) elsewhere, then \(\mathbb E_y f(x\star _Jy)=\sum _{S\subseteq J}\widehat f(S)\chi _S(x)\).
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube and \(J\subseteq [n]\). For uniform independent \(z\) and \(x\), and \(x\star _Jz\) keeping \(x\) on \(J\) and \(z\) elsewhere, \(\mathbb E_z\Pr _x(f(x)\ne f(x\star _Jz))=\tfrac 12\sum _{S\nsubseteq J}\widehat f(S)^2\).
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube and \(J\subseteq [n]\). There exists \(h:\{ 0,1\} ^n\to \{ -1,1\} \) depending only on \(J\) such that \(\Pr (f\ne h)\le \tfrac 12\sum _{S\nsubseteq J}\widehat f(S)^2\).
There exists an absolute \(B\ge 1\) such that for every linear threshold function \(f:\{ 0,1\} ^n\to \{ -1,1\} \) and \(0\lt \varepsilon \le 1/2\), there is \(k\in \mathbb N\) with \(k\le B/\varepsilon ^2\) and \(\sum _{|S|\gt k}\widehat f(S)^2\le \varepsilon \) under the uniform cube law.
There exists an absolute \(C\gt 0\) such that for every linear threshold function \(f:\{ 0,1\} ^n\to \{ -1,1\} \) under the uniform law and \(0\lt \varepsilon ,\eta \le 1/2\), there exist \(J\subseteq [n]\) and \(h:\{ 0,1\} ^n\to \{ -1,1\} \) depending only on \(J\) with \(|J|\le I(f)^{2+\eta }(1/\eta )^{C/\varepsilon ^2}\) and \(\Pr (f\ne h)\le \varepsilon \). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(n\in \mathbb N\) and \(-1\le \rho \le 1\), and put \(W(x,y)=2^{-n}\prod _{i=0}^{n-1}[(1+\rho (-1)^{x_i+y_i})/2]\) for cube points \(x,y\). Then \(W(x,y)\ge 0\), \(W(x,y)=W(y,x)\), \(\sum _yW(x,y)=2^{-n}\) for every \(x\), and \(\sum _{x,y}W(x,y)=1\).
Let \(A,B\subseteq \{ 0,1\} ^n\) have uniform volumes \(\alpha ,\beta \), and let real \(p,u,\rho \) satisfy \(1\le p\le u\), \(u\ge 2\), \(0\le \rho \le 1\), and \(\rho \le \sqrt{(p-1)/(u-1)}\). Then \(\mathbb E[\mathbf1_AT_\rho \mathbf1_B]\le \alpha ^{(u-1)/u}\beta ^{1/p}\).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\), \(i\in [n]\), and \(S\subseteq [n]\). For the signed coordinate derivative \(D_i f(x)=\tfrac 12(-1)^{x_i}(f(x)-f(x^{\oplus i}))\), where \(x^{\oplus i}\) flips coordinate \(i\), one has \(\widehat{D_i f}(S)=0\) if \(i\in S\), and \(\widehat{D_i f}(S)=\widehat f(S\cup \{ i\} )\) otherwise.
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) have Fourier degree at most \(k\), variance at most \(1\), and coordinate influences \(\operatorname {Inf}_i(f)\le \varepsilon \) for every \(i\), where \(\varepsilon \ge 0\). Then \(\sum _i\operatorname {Inf}_i(f)^2\le k\varepsilon \).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) have positive uniform variance, and let \(\rho \gt 0\). Writing \(I(f)=\sum _S|S|\widehat f(S)^2\), one has \(\operatorname {Var}(f)\rho ^{I(f)/\operatorname {Var}(f)}\le \sum _{S\subseteq [n]}|S|\rho ^{|S|}\widehat f(S)^2\).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) have Fourier degree at most \(k\in \mathbb N\), and let \(\rho \ge 1\). For the Fourier multiplier \(\widehat{T_\rho f}(S)=\rho ^{|S|}\widehat f(S)\), one has \(\| T_\rho f\| _2\le \rho ^k\| f\| _2\) under the uniform measure. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
For \(n\in \mathbb N\) and \(\rho \in \mathbb {R}\), the \(\rho \)-stable cube is the complete directed graph, including loops, on \(\{ 0,1\} ^n\), with weight \(w(x,y)=2^{-n}\prod _{i=1}^n\frac{1+\rho (-1)^{x_i+y_i}}{2}\). For \(-1\le \rho \le 1\), these weights are the joint probabilities of a uniform correlated pair.
For \(n\in \mathbb N\) and \(\delta \in \mathbb {R}\), the \(\delta \)-noisy cube is the \((1-2\delta )\)-stable cube: the complete directed graph on \(\{ 0,1\} ^n\) with weights \(2^{-n}\prod _{i=1}^n\frac{1+(1-2\delta )(-1)^{x_i+y_i}}{2}\).
Let \(p\in \mathbb {R}\) and \(f:\{ 0,1\} ^n\to \mathbb {R}\), define \(M_p(f)=0\) when \(p\le 0\) and \(f\) has a zero; otherwise set \(M_0(f)=\exp (\mathbb {E}\log |f|)\) and \(M_p(f)=(\mathbb {E}|f|^p)^{1/p}\) for \(p\ne 0\), with uniform expectation.
Let \(p\gt 0\) and \(f:\{ 0,1\} ^{n+1}\to \mathbb {R}\), and write \(M_{p,z}\) for the uniform power mean \((\mathbb {E}_z|\cdot |^p)^{1/p}\). Then \(M_{p,(x,b)}f(x,b)=M_{p,x}\bigl(M_{p,b}f(x,b)\bigr)\), where \(x\in \{ 0,1\} ^n\) and \(b\in \{ 0,1\} \).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) and \(\rho ,c\in \mathbb {R}\), the Fourier noise operator satisfies \(T_\rho (cf)=cT_\rho f\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(\rho \in \mathbb {R}\) and \(f:\{ 0,1\} ^0\to \mathbb {R}\), the Fourier noise operator satisfies \(T_\rho f=f\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(0\le \rho \le 1\) and \(f:\{ 0,1\} ^n\to [0,\infty )\). Then \(T_\rho f(x)\ge 0\) at every cube point \(x\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
For \(\rho ,a\in \mathbb {R}\), let \(f(b)=1+a(-1)^b\) on \(\{ 0,1\} \). Then \(T_\rho f(b)=1+\rho a(-1)^b\) for both bits \(b\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(f:\{ 0,1\} ^{n+1}\to \mathbb {R}\), \(\rho \in \mathbb {R}\), and \(f_b(z)=f(z,b)\). For \(x\in \{ 0,1\} ^n\) and \(b\in \{ 0,1\} \), \(T_\rho f(x,b)=T_\rho \bigl(t\mapsto T_\rho f_t(x)\bigr)(b)\); the outer operator acts on one bit and the inner operator on \(n\) bits. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(q\le p\le 1\), \(q\lt 1\), \(0\le \rho \le 1\), and \(\rho ^2\le (1-p)/(1-q)\). Every \(f:\{ 0,1\} ^n\to [0,\infty )\) satisfies \(M_q(T_\rho f)\ge M_p(f)\). For nonnegative functions \(u\), the extended mean is \(M_r(u)=(\mathbb {E}u^r)^{1/r}\) when \(r\gt 0\); the same formula applies for \(r\lt 0\) when \(u\) is positive everywhere. For positive \(u\), \(M_0(u)=\exp (\mathbb {E}\log u)\); if \(r\le 0\) and \(u\) has a zero, \(M_r(u)=0\). All expectations are uniform. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(q\le p\le 1\), \(q\lt 1\), \(0\le \rho \le 1\), and \(\rho ^2\le (1-p)/(1-q)\). Every \(f:\{ 0,1\} ^n\to [0,\infty )\) satisfies \(M_q(T_\rho f)\ge M_p(f)\). For nonnegative functions \(u\), the extended mean is \(M_r(u)=(\mathbb {E}u^r)^{1/r}\) when \(r\gt 0\); the same formula applies for \(r\lt 0\) when \(u\) is positive everywhere. For positive \(u\), \(M_0(u)=\exp (\mathbb {E}\log u)\); if \(r\le 0\) and \(u\) has a zero, \(M_r(u)=0\). All expectations are uniform. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(0\lt r\lt 1\) and \(u,v:\{ 0,1\} ^n\to [0,\infty )\). Then \(\mathbb {E}[uv]\ge M_r(u)M_{r/(r-1)}(v)\). For nonnegative functions \(u\), the extended mean is \(M_r(u)=(\mathbb {E}u^r)^{1/r}\) when \(r\gt 0\); the same formula applies for \(r\lt 0\) when \(u\) is positive everywhere. For positive \(u\), \(M_0(u)=\exp (\mathbb {E}\log u)\); if \(r\le 0\) and \(u\) has a zero, \(M_r(u)=0\). All expectations are uniform.
Let \(p\lt 1\), \(p\ne 0\), and \(f,g:\{ 0,1\} ^n\to [0,\infty )\). Then \(\mathbb {E}[fg]\ge M_p(f)M_{p/(p-1)}(g)\). For nonnegative functions \(u\), the extended mean is \(M_r(u)=(\mathbb {E}u^r)^{1/r}\) when \(r\gt 0\); the same formula applies for \(r\lt 0\) when \(u\) is positive everywhere. For positive \(u\), \(M_0(u)=\exp (\mathbb {E}\log u)\); if \(r\le 0\) and \(u\) has a zero, \(M_r(u)=0\). All expectations are uniform.
Let \(s\subseteq I\) and \(t\subseteq J\) be finite sets, let \(r\ge 1\), and let \(a:I\times J\to \mathbb {R}\) satisfy \(a(i,j)\ge 0\) for \(i\in s\), \(j\in t\). Then \((\sum _{j\in t}(\sum _{i\in s}a(i,j))^r)^{1/r}\le \sum _{i\in s}(\sum _{j\in t}a(i,j)^r)^{1/r}\).
Let \(0\lt q\lt p\) and \(0\le \rho \le 1\). Suppose \(M_q(T_\rho u)\ge M_p(u)\) for every nonnegative function \(u\) on one bit. Then \(M_q(T_\rho f)\ge M_p(f)\) for every \(n\in \mathbb N\) and \(f:\{ 0,1\} ^n\to [0,\infty )\), where \(M_r(u)=(\mathbb {E}|u|^r)^{1/r}\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(0\lt q\lt p\lt 1\), \(0\le \rho \le 1\), and \(\rho ^2=(1-p)/(1-q)\). Every \(f:\{ 0,1\} ^n\to [0,\infty )\) satisfies \(M_q(T_\rho f)\ge M_p(f)\), where \(M_r(u)=(\mathbb {E}|u|^r)^{1/r}\) under the uniform measure. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(q\lt 1\), \(0\le \rho \le \sigma \le 1\), and \(f:\{ 0,1\} ^n\to [0,\infty )\). Then \(M_q(T_\rho f)\ge M_q(T_\sigma f)\). For nonnegative functions \(u\), the extended mean is \(M_r(u)=(\mathbb {E}u^r)^{1/r}\) when \(r\gt 0\); the same formula applies for \(r\lt 0\) when \(u\) is positive everywhere. For positive \(u\), \(M_0(u)=\exp (\mathbb {E}\log u)\); if \(r\le 0\) and \(u\) has a zero, \(M_r(u)=0\). All expectations are uniform. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(a\le b\le 1\) and \(f:\{ 0,1\} ^n\to [0,\infty )\). Then \(M_a(f)\le M_b(f)\). For nonnegative functions \(u\), the extended mean is \(M_r(u)=(\mathbb {E}u^r)^{1/r}\) when \(r\gt 0\); the same formula applies for \(r\lt 0\) when \(u\) is positive everywhere. For positive \(u\), \(M_0(u)=\exp (\mathbb {E}\log u)\); if \(r\le 0\) and \(u\) has a zero, \(M_r(u)=0\). All expectations are uniform.
Let \(p,q\lt 1\), \(0\le \rho \le 1\), and \(\rho ^2\le (1-p)(1-q)\). For all \(f,g:\{ 0,1\} ^n\to [0,\infty )\), uniform expectation satisfies \(\mathbb {E}[fT_\rho g]\ge M_p(f)M_q(g)\). For nonnegative functions \(u\), the extended mean is \(M_r(u)=(\mathbb {E}u^r)^{1/r}\) when \(r\gt 0\); the same formula applies for \(r\lt 0\) when \(u\) is positive everywhere. For positive \(u\), \(M_0(u)=\exp (\mathbb {E}\log u)\); if \(r\le 0\) and \(u\) has a zero, \(M_r(u)=0\). All expectations are uniform. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(0\lt q\lt p\lt 1\), \(\rho \ge 0\), \(\rho ^2=(1-p)/(1-q)\), and integer \(K\ge 1\). If \(\binom p{2K}\le (p/q)\binom q{2K}\rho ^{2K}\), then \(\binom p{2K+2}\le (p/q)\binom q{2K+2}\rho ^{2K+2}\), with generalized binomial coefficients.
Let \(0\lt q\lt p\lt 1\), \(\rho \ge 0\), \(\rho ^2=(1-p)/(1-q)\), and \(|a|\lt 1\). For \(f(t)=1+a(-1)^t\) on \(\{ 0,1\} \), \(M_q(T_\rho f)\ge M_p(f)\), where \(M_r(u)=(\mathbb {E}|u|^r)^{1/r}\) under the uniform measure. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(0\lt q\lt p\lt 1\), \(\rho \ge 0\), and \(\rho ^2=(1-p)/(1-q)\). Every \(f:\{ 0,1\} \to [0,\infty )\) satisfies \(M_q(T_\rho f)\ge M_p(f)\), where \(M_r(u)=(\mathbb {E}|u|^r)^{1/r}\) under the uniform measure. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
For \(q,\lambda \in \mathbb {R}\), set \(u=\log ((1-\lambda )/\lambda )\) and define \(\rho (q,\lambda )=\sqrt{\sinh (u/q)/\sinh (u/(q/(q-1)))}\), where division by zero is zero, \(\log 0=0\), \(\log x=\log |x|\) for \(x\lt 0\), and \(\sqrt x=0\) for \(x\le 0\).
For \(\lambda \in \mathbb {R}\), define \(C(\lambda )=e\) if \(\lambda =1/2\), and \(C(\lambda )=((1-\lambda )/\lambda )^{1/(2(1-2\lambda ))}\) otherwise, where division by zero is zero; for \(x\lt 0\), \(x^r=|x|^r\cos (\pi r)\); and \(0^r=0\) for \(r\ne 0\), while \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(c\in \mathbb R\), and \(f:\Omega \to \mathbb R\). Then \(\operatorname {Stab}_{c^2}(f)=\| T_cf\| _{2,P}^2\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\operatorname {Stab}_a(h)=\mathbb E_P[hT_a h]\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), an event \(A\subseteq \Omega \), and \(p\gt 0\). Then \(\| \mathbf1_A\| _{p,P}^2=\Pr _P(A)^{2/p}\). For a real function \(h\) on \(\Omega \), use \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\Pr _P(A)=\mathbb E_P\mathbf1_A\), where \(\mathbf1_A\) is \(1\) on \(A\) and \(0\) elsewhere.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \le 1\), and let \(q\ge 2\) and \(0\le \rho \le \lambda ^{1-2/q}/(q-1)\). For every event \(A\subseteq \Omega \), \(\operatorname {Stab}_\rho (\mathbf1_A)\le \Pr _P(A)^{2-2/q}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\operatorname {Stab}_a(h)=\mathbb E_P[hT_a h]\); \(\Pr _P(A)=\mathbb E_P\mathbf1_A\), where \(\mathbf1_A\) is \(1\) on \(A\) and \(0\) elsewhere. The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(\rho \ge 0\), a coordinate \(i\), and \(f:\Omega \to \mathbb R\). Then \(\rho I_i^{(\rho )}(f)=\| T_{\sqrt\rho }L_if\| _{2,P}^2\), where \(I_i^{(\rho )}\) is stable influence and \(L_i\) is the coordinate Laplacian. For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(L_i h=h-E_{[n]\setminus \{ i\} }h\); \(I_i^{(a)}(h)=\sum _{S\ni i}a^{|S|-1}\mathbb E_Ph_S^2\) for \(a\in \mathbb R\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \{ -1,1\} \), and \(S\subseteq [n]\). For conditional expectation \(E_S\) onto \(S\), \(\mathbb E_P|f-E_Sf|=\mathbb E_P(f-E_Sf)^2\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \{ -1,1\} \), a coordinate \(i\), and \(1\le p\le 2\). Then \(\mathbb E_P|L_if|^p\le I_i(f)\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(L_i h=h-E_{[n]\setminus \{ i\} }h\); \(I_i(h)=\mathbb E_P(h-E_{[n]\setminus \{ i\} }h)^2\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(q\ge 2\), \(\rho \ge 0\), \(f:\Omega \to \{ -1,1\} \), and a coordinate \(i\). If \(\| T_{\sqrt\rho }L_if\| _{2,P}\le \| L_if\| _{q/(q-1),P}\), then \(\rho I_i^{(\rho )}(f)\le I_i(f)^{2-2/q}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(L_i h=h-E_{[n]\setminus \{ i\} }h\); \(I_i(h)=\mathbb E_P(h-E_{[n]\setminus \{ i\} }h)^2\); \(I_i^{(a)}(h)=\sum _{S\ni i}a^{|S|-1}\mathbb E_Ph_S^2\) for \(a\in \mathbb R\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \le 1\), and let \(q\ge 2\) and \(0\le \rho \le \lambda ^{1-2/q}/(q-1)\). For every \(f:\Omega \to \{ -1,1\} \) and coordinate \(i\), \(\rho I_i^{(\rho )}(f)\le I_i(f)^{2-2/q}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(I_i(h)=\mathbb E_P(h-E_{[n]\setminus \{ i\} }h)^2\); \(I_i^{(a)}(h)=\sum _{S\ni i}a^{|S|-1}\mathbb E_Ph_S^2\) for \(a\in \mathbb R\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \le 1\). For every \(f:\Omega \to \{ -1,1\} \) and coordinate \(i\), \(\sum _{S\ni i}(\sqrt\lambda /3)^{|S|}\mathbb E_Pf_S^2\le I_i(f)^{3/2}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(I_i(h)=\mathbb E_P(h-E_{[n]\setminus \{ i\} }h)^2\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \le 1/2\). If \(f:\Omega \to \{ -1,1\} \) has positive variance and \(a=I(f)/\operatorname {Var}_P(f)\), there is a coordinate \(i\) such that \((a^{-1})^2(9/\lambda )^{-a}\le I_i(f)\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(I_i(h)=\mathbb E_P(h-E_{[n]\setminus \{ i\} }h)^2\); \(I(h)=\sum _{i=0}^{n-1}I_i(h)=\sum _{S\subseteq [n]}|S|\mathbb E_Ph_S^2\); \(\operatorname {Var}_P(h)=\mathbb E_P(h-\mathbb E_Ph)^2\).
There is an absolute \(c\gt 0\) such that for every \(n\ge 2\), every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), with every coordinate atom at least \(0\lt \lambda \le 1/2\), and every \(f:\Omega \to \{ -1,1\} \), some coordinate \(i\) satisfies \(c\operatorname {Var}_P(f)\log n/(n\log (1/\lambda ))\le I_i(f)\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(I_i(h)=\mathbb E_P(h-E_{[n]\setminus \{ i\} }h)^2\); \(\operatorname {Var}_P(h)=\mathbb E_P(h-\mathbb E_Ph)^2\).
There is an absolute \(C\gt 0\) such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), with every coordinate atom at least \(0\lt \lambda \le 1/2\), every \(f:\Omega \to \{ -1,1\} \), and every \(0\lt \varepsilon \le 1\), there exist \(J\subseteq [n]\) and \(g:\Omega \to \{ -1,1\} \) depending only on \(J\) with \(|J|\le (1/\lambda )^{CI(f)/\varepsilon }\) and \(\Pr _P(f\ne g)\le \varepsilon \). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(I_i(h)=\mathbb E_P(h-E_{[n]\setminus \{ i\} }h)^2\); \(I(h)=\sum _{i=0}^{n-1}I_i(h)=\sum _{S\subseteq [n]}|S|\mathbb E_Ph_S^2\); \(\Pr _P(A)=\mathbb E_P\mathbf1_A\), where \(\mathbf1_A\) is \(1\) on \(A\) and \(0\) elsewhere.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \lt 1/2\), and let \(q\gt 2\). Define \(r_q(\lambda )=\sqrt{\frac{\sinh (\log ((1-\lambda )/\lambda )/q)}{\sinh (\log ((1-\lambda )/\lambda )/(q/(q-1)))}}\); if \(0\le \rho \le r_q(\lambda )^2\), then \(\operatorname {Stab}_\rho (\mathbf1_A)\le \Pr _P(A)^{2-2/q}\) for every event \(A\subseteq \Omega \). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\operatorname {Stab}_a(h)=\mathbb E_P[hT_a h]\); \(\Pr _P(A)=\mathbb E_P\mathbf1_A\), where \(\mathbf1_A\) is \(1\) on \(A\) and \(0\) elsewhere. The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \lt 1/2\), and let \(q\gt 2\). Define \(r_q(\lambda )=\sqrt{\frac{\sinh (\log ((1-\lambda )/\lambda )/q)}{\sinh (\log ((1-\lambda )/\lambda )/(q/(q-1)))}}\); if \(0\le \rho \le r_q(\lambda )^2\), then \(\rho I_i^{(\rho )}(f)\le I_i(f)^{2-2/q}\) for every \(f:\Omega \to \{ -1,1\} \) and coordinate \(i\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(I_i(h)=\mathbb E_P(h-E_{[n]\setminus \{ i\} }h)^2\); \(I_i^{(a)}(h)=\sum _{S\ni i}a^{|S|-1}\mathbb E_Ph_S^2\) for \(a\in \mathbb R\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(0\le c\le 1\), and \(f:\Omega \to \mathbb R\). Then \(\| T_cf\| _{2,P}\le \| f\| _{2,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(c\ge 1\), \(k\in \mathbb N\), and \(f:\Omega \to \mathbb R\) of degree at most \(k\). Then \(\| T_cf\| _{2,P}\le c^k\| f\| _{2,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); degree at most \(k\) means \(h_S=0\) for every \(|S|\gt k\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(\lambda \), where \(0\lt \lambda \le 1\), and let \(q\gt 2\). For \(k\in \mathbb N\) and \(f:\Omega \to \mathbb R\) of degree at most \(k\), \(\| f\| _{q,P}\le (\sqrt{q-1}\lambda ^{1/q-1/2})^k\| f\| _{2,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); degree at most \(k\) means \(h_S=0\) for every \(|S|\gt k\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \lt 1/2\). For \(f:\Omega \to \mathbb R\) of degree at most \(k\in \mathbb N\), \(\| f\| _{2,P}\le C_\lambda ^k\| f\| _{1,P}\), where \(C_\lambda =((1-\lambda )/\lambda )^{1/(2(1-2\lambda ))}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); degree at most \(k\) means \(h_S=0\) for every \(|S|\gt k\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(\lambda \), where \(0\lt \lambda \le 1/2\). Define \(C_\lambda =\begin{cases} e,& \lambda =1/2,\\ ((1-\lambda )/\lambda )^{1/(2(1-2\lambda ))},& \lambda \ne 1/2.\end{cases}\) For \(f:\Omega \to \mathbb R\) of degree at most \(k\in \mathbb N\), \(\| f\| _{2,P}\le C_\lambda ^k\| f\| _{1,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); degree at most \(k\) means \(h_S=0\) for every \(|S|\gt k\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \le 1/2\). If \(f:\Omega \to \mathbb R\) has degree at most \(k\in \mathbb N\) and positive variance, then \(\tfrac 14(e^2/(2\lambda ))^{-k}\le \Pr _P(\mathbb E_Pf\lt f)\) and \((15/\lambda )^{-k}\le \tfrac 14(e^2/(2\lambda ))^{-k}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\operatorname {Var}_P(h)=\mathbb E_P(h-\mathbb E_Ph)^2\); \(\Pr _P(A)=\mathbb E_P\mathbf1_A\), where \(\mathbf1_A\) is \(1\) on \(A\) and \(0\) elsewhere; degree at most \(k\) means \(h_S=0\) for every \(|S|\gt k\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \le 1\). For \(q\ge 2\) and \(f:\Omega \to \mathbb R\) of degree at most \(k\in \mathbb N\), \(\| f\| _{q,P}\le (\sqrt q\lambda ^{1/q-1/2})^k\| f\| _{2,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); degree at most \(k\) means \(h_S=0\) for every \(|S|\gt k\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(0\lt \lambda \le 1\). If \(k\gt 0\) is an integer, \(f:\Omega \to \mathbb R\) has degree at most \(k\), \(\| f\| _{2,P}\gt 0\), and \(t\ge (\sqrt{2e/\lambda })^k\), then \(\Pr _P(|f|\ge t\| f\| _{2,P})\le \lambda ^k\exp (-k\lambda t^{2/k}/(2e))\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\Pr _P(A)=\mathbb E_P\mathbf1_A\), where \(\mathbf1_A\) is \(1\) on \(A\) and \(0\) elsewhere; degree at most \(k\) means \(h_S=0\) for every \(|S|\gt k\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(0\le c\le 1\), and \(f:\Omega \to \mathbb R\). Then \(\| T_cf\| _{2,P}\le \| f\| _{2,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
For every \(n\in \mathbb N\), a finite probability product consists of finite sets \(\Omega _i\) for \(i\in \{ 0,\ldots ,n-1\} \) and real weights \(w_i(a)\gt 0\) satisfying \(\sum _{a\in \Omega _i}w_i(a)=1\). The coordinate sets and weights may depend on \(i\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), with product masses \(m(x)\). For \(f:\prod _i\Omega _i\to \mathbb R\), define \(\mathbb E_Pf=\sum _xm(x)f(x)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) and let \(A\) be an event in its outcome space. Its indicator \(\mathbf1_A\) equals \(1\) on \(A\) and \(0\) elsewhere.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(q\in \mathbb R\), and \(f:\prod _i\Omega _i\to \mathbb R\). Define \(\| f\| _{q,P}=(\mathbb E_P|f|^q)^{1/q}\) using real powers and division; for \(q\gt 0\) this is the usual moment expression.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S\subseteq \{ 0,\ldots ,n-1\} \), and \(f:\prod _i\Omega _i\to \mathbb R\). Say that \(f\) depends only on \(S\) if \(f(x)=f(y)\) whenever \(x_i=y_i\) for every \(i\in S\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S\subseteq \{ 0,\ldots ,n-1\} \), and \(f:\prod _i\Omega _i\to \mathbb R\). Define \(E_Sf(x)=\mathbb E_{Y\sim P}f(z)\), where \(z_i=x_i\) for \(i\in S\) and \(z_i=Y_i\) otherwise.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S\subseteq \{ 0,\ldots ,n-1\} \), and \(f:\prod _i\Omega _i\to \mathbb R\). Its \(S\)-component is \(f_S=\sum _{T\subseteq S}(-1)^{|S|-|T|}E_Tf\), where \(E_T\) conditions on the coordinates in \(T\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(k\in \mathbb N\), and \(f:\prod _i\Omega _i\to \mathbb R\). The predicate that \(f\) has degree at most \(k\) means that every orthogonal component \(f_S\) with \(|S|\gt k\) vanishes identically.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(k\in \mathbb N\), and \(f:\prod _i\Omega _i\to \mathbb R\). Define \(f^{\le k}(x)=\sum _{|S|\le k}f_S(x)\), where \(f_S\) is the orthogonal component indexed by \(S\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(\rho \in \mathbb R\), and \(f:\prod _i\Omega _i\to \mathbb R\). Define \(T_\rho f(x)=\sum _S\rho ^{|S|}f_S(x)\), where \(f_S\) are the orthogonal components and the powers have natural exponents.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), a coordinate \(i\), and \(f:\prod _j\Omega _j\to \mathbb R\). Define \(L_if=f-E_{[n]\setminus \{ i\} }f\), where \(E_S\) conditions on the coordinates in \(S\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), a coordinate \(i\), and \(f:\prod _j\Omega _j\to \mathbb R\). Define \(I_i(f)=\mathbb E_P(L_if)^2\), where \(L_if\) subtracts the conditional expectation over coordinate \(i\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(\rho \in \mathbb R\), a coordinate \(i\), and \(f:\prod _j\Omega _j\to \mathbb R\). Define \(I_i^{(\rho )}(f)=\sum _{S\ni i}\rho ^{|S|-1}\mathbb E_Pf_S^2\), with natural exponents, including \(\rho =0\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(\rho \in \mathbb R\), and \(f:\prod _i\Omega _i\to \mathbb R\). Define \(\operatorname {Stab}_\rho (f)=\mathbb E_P[fT_\rho f]\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with coordinate weights \(w_i\) and let \(\lambda \in \mathbb R\). The atom lower bound \(\lambda \) means that \(\lambda \le w_i(a)\) for every coordinate \(i\) and outcome \(a\in \Omega _i\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) and \(f:\prod _i\Omega _i\to \mathbb R\). The Boolean-valued predicate means \(f(x)\in \{ -1,1\} \) at every outcome \(x\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S\subseteq [n]\), and \(F:\Omega \times \Omega \to \mathbb R\), where \(\Omega =\prod _i\Omega _i\). For independent \(X,Y\sim P\), let \(X\mathbin {\star _S}Y\) keep \(X_i\) on \(S\) and \(Y_i\) elsewhere; then \(\mathbb E F(X\mathbin {\star _S}Y,Y\mathbin {\star _S}X)=\mathbb EF(X,Y)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), let \(s\) be a finite index set, and let \(a_j\in \mathbb R\) and \(F_j:\Omega \to \mathbb R\). Then \(\mathbb E_P\sum _{j\in s}a_jF_j=\sum _{j\in s}a_j\mathbb E_PF_j\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S,T\subseteq [n]\), \(f:\Omega \to \mathbb R\), and \(x\in \Omega \). For coordinate conditional expectations, \(E_S(E_Tf)(x)=E_{S\cap T}f(x)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S\subseteq [n]\), and \(f,g:\Omega \to \mathbb R\). For conditional expectation \(E_S\) onto \(S\), \(\mathbb E_P[fE_Sg]=\mathbb E_P[(E_Sf)g]\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), a finite index set \(s\), and functions \(F_j:\Omega \to \mathbb R\). Then \(\mathbb E_P\sum _{j\in s}F_j=\sum _{j\in s}\mathbb E_PF_j\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) and \(f:\Omega \to \mathbb R\) with orthogonal components \(f_S\). Then \(\mathbb E_Pf^2=\sum _{S\subseteq [n]}\mathbb E_Pf_S^2\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(J\subseteq [n]\), and functions \(f,h:\Omega \to \mathbb R\). If \(f\) takes values in \(\{ -1,1\} \) and \(h\) depends only on \(J\), there exists \(g:\Omega \to \{ -1,1\} \) depending only on \(J\) such that \(\Pr _P(f\ne g)\le \mathbb E_P(f-h)^2\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\) with variance \(V\gt 0\), and \(\rho \gt 0\). Writing \(I=\sum _iI_i(f)\) and \(f_S\) for the orthogonal components, \(V\rho ^{I/V}\le \sum _S|S|\rho ^{|S|}\mathbb E_Pf_S^2\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(q\gt 2\), \(c\in \mathbb R\), and \(f:\Omega \to \mathbb R\). If \(\| f\| _{q,P}\le e^c\| f\| _{2,P}\), then \(\| f\| _{2,P}\le e^{cq/(q-2)}\| f\| _{1,P}\).
Let \(0\lt \lambda \lt 1/2\), define \(r_q(\lambda )=\sqrt{\frac{\sinh (\log ((1-\lambda )/\lambda )/q)}{\sinh (\log ((1-\lambda )/\lambda )/(q/(q-1)))}}\), and put \(C_\lambda =((1-\lambda )/\lambda )^{1/(2(1-2\lambda ))}\). Then \(\lim _{q\to 2^+}\exp (-q\log r_q(\lambda )/(q-2))=C_\lambda \).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) and let \(g:\Omega \to \mathbb R\) have mean zero and positive second moment. If \(C\gt 0\) and \(\| g\| _{2,P}\le C\| g\| _{1,P}\), then \(1/(4C^2)\le \Pr _P(g\gt 0)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), \(J\subseteq [n]\), \(k\in \mathbb N\), \(0\le \rho \le 1\), and \(\tau \ge 0\). If \(I_i(f)\le \tau \) for \(i\notin J\) and \(\sum _{S\ni i}\rho ^{|S|}\mathbb E_Pf_S^2\le I_i(f)^{3/2}\) for every \(i\), then \(\rho ^k\sum _{|S|\le k,\ S\nsubseteq J}\mathbb E_Pf_S^2\le \sqrt\tau \, I(f)\).
Let \(0\lt \lambda \le 1/2\), \(I,\varepsilon \gt 0\), and \(k\in \mathbb N\). Put \(\rho =\sqrt\lambda /3\) and \(\tau =(\varepsilon /(2I))^2\rho ^{2k}\); then \(0\lt \rho \le 1\), \(\tau \gt 0\), \(\sqrt\tau I=(\varepsilon /2)\rho ^k\), and \(I/\tau =(4I^3/\varepsilon ^2)(9/\lambda )^k\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(p\ge 1\), \(q\gt 1\), and \(0\le \rho \le 1\). If every coordinate probability law satisfies \(\| \rho g+(1-\rho )\mathbb E_i g\| _{q,i}\le \| g\| _{p,i}\) for every real \(g\) on that coordinate, then \(\| T_\rho f\| _{q,P}\le \| f\| _{p,P}\) for every \(f:\Omega \to \mathbb R\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\| g\| _{p,i}=(\sum _{a\in \Omega _i}w_i(a)|g(a)|^p)^{1/p}\) and \(\mathbb E_i g=\sum _{a\in \Omega _i}w_i(a)g(a)\), where \(w_i\) gives the probability masses in coordinate \(i\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(1\le p\le q\), and \(0\le \rho \le 1\). If every coordinate satisfies \(\| \rho g+(1-\rho )\mathbb E_i g\| _{q,i}\le \| g\| _{p,i}\) for every real coordinate function \(g\), then \(\| T_\rho f\| _{q,P}\le \| f\| _{p,P}\) for every \(f:\Omega \to \mathbb R\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\| g\| _{p,i}=(\sum _{a\in \Omega _i}w_i(a)|g(a)|^p)^{1/p}\) and \(\mathbb E_i g=\sum _{a\in \Omega _i}w_i(a)g(a)\), where \(w_i\) gives the probability masses in coordinate \(i\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(I\) be a finite set with nonnegative weights \(w_i\) satisfying \(\sum _iw_i=1\), and let \(g:I\to \mathbb R\) and \(0\le \rho \le \sigma \). Then \(\rho ^2\sum _iw_ig_i^2+(1-\rho ^2)(\sum _iw_ig_i)^2\le \sigma ^2\sum _iw_ig_i^2+(1-\sigma ^2)(\sum _iw_ig_i)^2\).
Let \(I\) be finite with weights \(w_i\ge \lambda \), \(\sum _iw_i=1\), \(0\lt \lambda \lt 1/2\), and \(q\gt 2\). Put \(r=\sqrt{\sinh (\log ((1-\lambda )/\lambda )/q)/\sinh (\log ((1-\lambda )/\lambda )/(q/(q-1)))}\). For \(0\le \rho \le r\), the operator \(Ng=\rho g+(1-\rho )\sum _iw_ig_i\) satisfies \(\| Ng\| _{q,w}\le \| g\| _{2,w}\) and \(\| Ng\| _{2,w}\le \| g\| _{q/(q-1),w}\). Here \(\| h\| _{p,w}=(\sum _{i\in I}w_i|h_i|^p)^{1/p}\) for \(p\gt 0\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(\lambda \), where \(0\lt \lambda \lt 1/2\), and let \(q\gt 2\). Put \(r=\sqrt{\sinh (\log ((1-\lambda )/\lambda )/q)/\sinh (\log ((1-\lambda )/\lambda )/(q/(q-1)))}\). For \(0\le \rho \le r\) and every \(f:\Omega \to \mathbb R\), \(\| T_\rho f\| _{q,P}\le \| f\| _{2,P}\) and \(\| T_\rho f\| _{2,P}\le \| f\| _{q/(q-1),P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(1\), and let \(q\gt 2\) and \(\rho \in \mathbb R\). For every \(f:\Omega \to \mathbb R\), \(\| T_\rho f\| _{q,P}\le \| f\| _{2,P}\) and \(\| T_\rho f\| _{2,P}\le \| f\| _{q/(q-1),P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with every coordinate atom at least \(\lambda \), where \(0\lt \lambda \le 1\), and let \(q\gt 2\). For \(0\le \rho \le \lambda ^{1/2-1/q}/\sqrt{q-1}\) and every \(f:\Omega \to \mathbb R\), \(\| T_\rho f\| _{q,P}\le \| f\| _{2,P}\) and \(\| T_\rho f\| _{2,P}\le \| f\| _{q/(q-1),P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a probability law on a finite outcome space \(\Omega \), with every atom of mass at least \(\lambda \gt 0\), where \(0\lt \lambda \le 1\), and let \(q\gt 2\). For \(0\le \rho \le \lambda ^{1/2-1/q}/\sqrt{q-1}\), the operator \(T_\rho f=\rho f+(1-\rho )\mathbb E_Pf\) satisfies \(\| T_\rho f\| _{q,P}\le \| f\| _{2,P}\) and \(\| T_\rho f\| _{2,P}\le \| f\| _{q/(q-1),P}\) for every real \(f\). For a real function \(h\) on \(\Omega \), use \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with coordinate weights \(w_i\), \(S\subseteq [n]\), and \(f:\Omega \to \mathbb R\). Then \(E_Sf(x)=\sum _y[\prod _{i\in S}\mathbf1_{x_i=y_i}\prod _{i\notin S}w_i(y_i)]f(y)\) for every \(x\in \Omega \).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with coordinate weights \(w_i\), \(S\subseteq [n]\), and \(f:\Omega \to \mathbb R\). Its orthogonal component satisfies \(f_S(x)=\sum _y[\prod _{i\in S}(\mathbf1_{x_i=y_i}-w_i(y_i))\prod _{i\notin S}w_i(y_i)]f(y)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) with coordinate weights \(w_i\), \(\rho \in \mathbb R\), and \(f:\Omega \to \mathbb R\). Then \(T_\rho f(x)=\sum _y[\prod _i(\rho \mathbf1_{x_i=y_i}+(1-\rho )w_i(y_i))]f(y)\) for every \(x\in \Omega \).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(0\le \rho \le 1\), and \(f:\Omega \to \mathbb R\). Then \(\| T_\rho f\| _{1,P}\le \| f\| _{1,P}\).
Let \(I\) be a finite set with nonnegative weights \(w_i\) and define \(\| f\| _{r,w}=(\sum _iw_i|f_i|^r)^{1/r}\). If an operator \(T\) satisfies \(\sum _iw_i(Tf)_ig_i=\sum _iw_if_i(Tg)_i\) for all \(f,g\) and \(\| Tg\| _{q/(q-1),w}\le \| g\| _{q/(q-1),w}\) for every \(g\), where \(q\gt 1\), then \(\| Tf\| _{q,w}\le \| f\| _{q,w}\) for every \(f\).
Let \(\alpha \) be finite, let \(E(S)\ge 0\) for every \(S\subseteq \alpha \), let \(J\subseteq \alpha \), \(k\in \mathbb N\), and \(0\le \rho \le 1\). Then \(\rho ^k\sum _{|S|\le k,\ S\nsubseteq J}E(S)\le \sum _{i\notin J}\sum _{S\ni i}\rho ^{|S|}E(S)\).
Let \(\Omega _i\) be finite sets, \(w_i(a)\ge 0\) weights, and \(k_i(a,b)\ge 0\) kernels for \(i\in [n]\). For \(p,r\ge 1\), suppose \(\sum _{a,b}k_i(a,b)u(a)v(b)\le \| u\| _{p,w_i}\| v\| _{r,w_i}\) for every coordinate \(i\) and real \(u,v\). Then \(\sum _{x,y}(\prod _i k_i(x_i,y_i))f(x)g(y)\le \| f\| _{p,w}\| g\| _{r,w}\) for all real \(f,g\) on \(\prod _i\Omega _i\), where \(w(x)=\prod _iw_i(x_i)\) and \(\| h\| _{s,w}=(\sum _xw(x)|h(x)|^s)^{1/s}\).
Let \(\Omega _i\) be finite sets, \(w_i(a)\ge 0\) weights, and \(k_i(y,x)\ge 0\) kernels for \(i\in [n]\). For \(p\ge 1\) and \(q\gt 1\), suppose \(\| \sum _xk_i(\, \cdot \, ,x)h(x)\| _{q,w_i}\le \| h\| _{p,w_i}\) for every coordinate \(i\) and real \(h\). Then the product operator \(Tf(y)=\sum _x\prod _i k_i(y_i,x_i)f(x)\) satisfies \(\| Tf\| _{q,w}\le \| f\| _{p,w}\) for every real \(f\) on \(\prod _i\Omega _i\), with \(w(x)=\prod _iw_i(x_i)\) and \(\| h\| _{s,w}=(\sum _xw(x)|h(x)|^s)^{1/s}\).
For a real function \(X\) on a measure space \((\Omega ,\mu )\) and extended exponents \(p,q\in [0,\infty ]\), \((p,q,\rho )\)-hypercontractivity means \(1\le p\le q\), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for every \(a,b\in \mathbb {R}\), with extended-valued norms and the essential supremum at infinity.
For a real function \(X\) on a measure space \((\Omega ,\mu )\) and \(q\in \mathbb {R}\), define \(N_q(X)\) by taking the extended norm at exponent \(\max (q,0)\) and converting it to a real number; a finite value is retained and an infinite value is assigned \(0\).
A real function \(X\) on \((\Omega ,\mu )\) has atom lower bound \(\lambda \in \mathbb {R}\) when there exists a finite \(s\subseteq \mathbb {R}\) supporting \(X\) almost everywhere and every \(x\in s\) satisfies \(\mu (X=x)\ge \max (\lambda ,0)\).
Let \(X\) be centered on a probability space, let \(q\gt 2\), and suppose \(H_{2,q,\rho }(X)\). Then \(H_{q/(q-1),2,\rho }(X)\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(X\) be a real function on an arbitrary measure space, suppose \(H_{p,q,\rho }(X)\), and let \(c\in \mathbb {R}\). Then \(H_{p,q,\rho }(cX)\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(X\) be centered on a probability space, suppose \(H_{p,q,\rho }(X)\), and let \(0\le \sigma \le \rho \). Then \(H_{p,q,\sigma }(X)\), including extended exponents equal to infinity. Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(\mu \) be a finite measure and \(X:\Omega \to \mathbb {R}\) measurable. If a finite set supports \(X\) almost everywhere and each listed value has mass at least \(\max (\lambda ,0)\) for some real \(\lambda \), then \(X\in L^p(\mu )\) for every \(p\in [0,\infty ]\), including essential boundedness.
Let \(X_1,\ldots ,X_n\) be mutually independent real random variables on a probability space, each belonging to \(L^4\). For arbitrary real coefficients \(F(S)\), the multilinear evaluation \(\sum _{S\subseteq [n]}F(S)\prod _{i\in S}X_i\) belongs to \(L^4\).
Let \(s\) be a finite subset of an index set \(I\), let \(i\notin s\), let \(F\) assign real coefficients to finite subsets of \(I\), and let \(X_j:\Omega \to \mathbb {R}\). For every \(\omega \in \Omega \), \(\sum _{T\subseteq s\cup \{ i\} }F(T)\prod _{j\in T}X_j(\omega )=\sum _{T\subseteq s}F(T)\prod _{j\in T}X_j(\omega )+X_i(\omega )\sum _{T\subseteq s}F(T\cup \{ i\} )\prod _{j\in T}X_j(\omega )\).
Let \(q\gt 2\) and \(\lambda \gt 0\). Let \(X\) be a measurable real random variable on a probability space, supported almost surely on a finite set whose listed atoms all have probability at least \(\lambda \). Then \(\| X\| _q\le (1/\lambda )^{1/2-1/q}\| X\| _2\).
Let \(q\gt 2\), \(0\lt \lambda \le 1\), and let \(X\) be measurable and centered on a probability space, supported almost surely on a finite set whose atoms each have mass at least \(\lambda \). Setting \(\rho =\lambda ^{1/2-1/q}/(2\sqrt{q-1})\), both \(H_{2,q,\rho }(X)\) and \(H_{q/(q-1),2,\rho }(X)\) hold. Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(q\gt 2\), \(0\lt \lambda \le 1\), and let measurable \(X\) have a symmetric law on a probability space, supported almost surely on a finite set whose atoms each have mass at least \(\lambda \). Setting \(\rho =\lambda ^{1/2-1/q}/\sqrt{q-1}\), both \(H_{2,q,\rho }(X)\) and \(H_{q/(q-1),2,\rho }(X)\) hold. Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(q\gt 2\) and \(0\lt \lambda \lt 1/2\), and let \(X\) be measurable and centered on a probability space, supported almost surely on a finite set whose atoms each have probability at least \(\lambda \). Both \(H_{2,q,\rho _*}(X)\) and \(H_{q/(q-1),2,\rho _*}(X)\) hold. Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(q\gt 2\), \(0\lt \lambda \lt 1/2\), and \(0\le \rho \lt 1\). Let measurable \(X\) on a probability space satisfy \(\Pr [X=1-\lambda ]=\lambda \) and \(\Pr [X=-\lambda ]=1-\lambda \). Then \(H_{2,q,\rho }(X)\) holds if and only if \(\rho \le \rho _*(q,\lambda )\), and the same equivalence holds for \(H_{q/(q-1),2,\rho }(X)\). Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(\mu \) be a finite measure, let \(X:\Omega \to \mathbb {R}\) be measurable and supported almost everywhere on a finite set \(s\subseteq \mathbb {R}\), and let \(f:\mathbb {R}\to \mathbb {R}\) be arbitrary. Then \(\int f(X)\, d\mu =\sum _{x\in s}\mu (X=x)f(x)\).
Let \(X\) be measurable on a probability space, let \(x\ne y\) and \(0\le \lambda \le 1\), and suppose \(\Pr [X=x]=\lambda \) and \(\Pr [X=y]=1-\lambda \). For \(q\gt 0\) and \(a,b\in \mathbb {R}\), \(\| a+bX\| _q^2=(\lambda |a+bx|^q+(1-\lambda )|a+by|^q)^{2/q}\).
Let \(X\) be measurable on a probability space and supported almost surely on a finite set \(s\subseteq \mathbb {R}\). If \(\lambda \ge 0\) and \(\Pr [X=x]\ge \lambda \) for every \(x\in s\), then the real atom weights \(w_x=\Pr [X=x]\) satisfy \(\sum _{x\in s}w_x=1\) and \(w_x\ge \lambda \) for every \(x\in s\).
Let \(X,Y\) be independent real random variables on a probability space. If both satisfy \(H_{p,q,\rho }\), then \(X+Y\) also satisfies \(H_{p,q,\rho }\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(q\gt 2\) and let \(X\in L^q\) be centered on a probability space, with \(\| X\| _2=1\) and \(\| X\| _q=C\gt 0\). Then \(H_{2,q,\rho }(X)\) holds for \(\rho =1/(2C\sqrt{q-1})\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \((X_i)_{i\in I}\) be mutually independent real random variables on a probability space, with \(I\) finite, \(X_i\in L^2\), \(\mathbb {E}X_i=0\), and \(\mathbb {E}X_i^2=1\) for every \(i\). For \(S,T\subseteq I\), \(\mathbb {E}[(\prod _{i\in S}X_i)(\prod _{i\in T}X_i)]=1\) if \(S=T\), and \(0\) otherwise.
Let \(X_1,\ldots ,X_n\) be mutually independent real random variables on a probability space, each in \(L^4\), centered, and with second moment \(1\). For real coefficient families \(F,G\) indexed by subsets of \([n]\), writing \(F(X)=\sum _S F(S)\prod _{i\in S}X_i\) and similarly for \(G\), one has \(\mathbb {E}[F(X)G(X)]=\sum _S F(S)G(S)\).
Let \(X_1,\ldots ,X_n\) be mutually independent real random variables on a probability space, each in \(L^4\), centered, and with second moment \(1\). For \(F(X)=\sum _{S\subseteq [n]}F(S)\prod _{i\in S}X_i\) with real coefficients, \(\mathbb {E}[F(X)^2]=\sum _S F(S)^2\).
Let \(X_1,\ldots ,X_n\) be mutually independent real random variables on a probability space, each in \(L^4\), centered, and with second moment \(1\). For arbitrary real coefficients \(F(S)\), \(\mathbb {E}[\sum _{S\subseteq [n]}F(S)\prod _{i\in S}X_i]=F(\varnothing )\).
Let \(F\) be a real coefficient family indexed by subsets of \([n]\), let \(X_i:\Omega \to \mathbb {R}\), and let \(s\subseteq [n]\). If \(F(T)=0\) whenever \(T\nsubseteq s\), then for every \(\omega \), \(\sum _{T\subseteq [n]}F(T)\prod _{j\in T}X_j(\omega )=\sum _{T\subseteq s}F(T)\prod _{j\in T}X_j(\omega )\).
Let \(F\) be real multilinear coefficients on \(n\) inputs and let \(i\in [n]\) satisfy \(F(S)=0\) for every \(S\) containing \(i\). If functions \(X_j,Y_j:\Omega \to \mathbb {R}\) agree for every \(j\ne i\), then \(\sum _S F(S)\prod _{j\in S}X_j=\sum _S F(S)\prod _{j\in S}Y_j\) as functions on \(\Omega \).
Let \(X,X^{\prime }\) be independent identically distributed real random variables on a probability space, let \(1\le q\le \infty \), assume \(X\in L^q\) and \(\mathbb {E}X=0\), and let \(a\in \mathbb {R}\). Then \(\| a+X\| _q\le \| a+X-X^{\prime }\| _q\), with the essential supremum at \(q=\infty \).
Let \(X\) be a centered real random variable on a probability space, let \(1\le q\le \infty \), and assume \(X\in L^q\). If \(r\) is an independent measurable uniform sign in \(\{ -1,1\} \), then for every \(a\in \mathbb {R}\), \(\| a+X/2\| _q\le \| a+rX\| _q\), including the essential-supremum norm.
On a probability space let real \(A,D,X\in L^4\), assume \(X\) is independent of \((A,D)\), and suppose \(\mathbb {E}X=\mathbb {E}X^3=0\). Then \(\mathbb {E}(A+XD)^2=\mathbb {E}A^2+\mathbb {E}X^2\mathbb {E}D^2\) and \(\mathbb {E}(A+XD)^4=\mathbb {E}A^4+6\mathbb {E}X^2\mathbb {E}[A^2D^2]+\mathbb {E}X^4\mathbb {E}D^4\).
Let \((X_j)_{j\in [n]}\) be mutually independent, almost everywhere measurable real functions on a measure space \((\Omega ,\mu )\), let \(s\subseteq [n]\) and \(i\notin s\), and let \(F,G\) be real coefficient families on subsets of \([n]\). Then \(X_i\) is independent of the pair \((\sum _{T\subseteq s}F(T)\prod _{j\in T}X_j,\sum _{T\subseteq s}G(T)\prod _{j\in T}X_j)\).
Let \(X_1,\ldots ,X_n\in L^4\) be mutually independent real random variables on a probability space, with \(\mathbb {E}X_i=\mathbb {E}X_i^3=0\) and \(\mathbb {E}X_i^4\le B(\mathbb {E}X_i^2)^2\) for every \(i\), where \(B\ge 1\). If \(F(X)=\sum _{S\subseteq [n]}F(S)\prod _{i\in S}X_i\) has coefficients zero whenever \(|S|\gt k\), then \(F(X)\in L^4\) and \(\mathbb {E}F(X)^4\le \max (B,9)^k(\mathbb {E}F(X)^2)^2\).
Let \(q\gt 2\), \(0\lt \lambda \lt 1/2\), and let \(X\) be a centered measurable real random variable on a probability space, supported almost surely on a finite set whose atoms each have mass at least \(\lambda \). Then \(H_{2,q,\rho _*(q,\lambda )}(X)\) holds. Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(I\) be finite, let real weights satisfy \(\sum _iw_i=1\), and let \(f:I\to \mathbb {R}\) take values in \(\{ a,b\} \), where \(a,b\in \mathbb {R}\). For any \(\varphi :\mathbb {R}\to \mathbb {R}\), setting \(\gamma =\sum _{i:f_i=a}w_i\) gives \(\sum _iw_i\varphi (f_i)=\gamma \varphi (a)+(1-\gamma )\varphi (b)\).
Let \(I\) be finite, \(w_i\gt 0\), \(\sum _iw_i=1\), and \(0\lt \lambda \le w_i\) for every \(i\). If \(f:I\to \mathbb {R}\) takes values in \(\{ a,b\} \), then either \(f\) is identically \(a\), or identically \(b\), or \(\lambda \le \sum _{i:f_i=a}w_i\le 1-\lambda \).
Let \(I\) be finite, \(w_i\gt 0\), \(p\gt 0\), and \(\rho \in \mathbb {R}\). Define \(G(f)=\sum _iw_i|f_i|^p\) and \(F(f)=\rho ^2\sum _iw_if_i^2+(1-\rho ^2)(\sum _iw_if_i)^2\). If \(G(f)=1\) implies \(F(f)\le 1\) for every real vector \(f\), then \(F(f)\le G(f)^{2/p}\) for every \(f\).
Let \(I\) be finite, let \(\sum _iw_i=1\), let \(q\gt 2\), \(0\lt \lambda \lt 1/2\), and assume \(w_i\ge \lambda \) for all \(i\). Put \(p=q/(q-1)\) and \(\rho =\rho _*(q,\lambda )\). Every real vector \(f\) satisfies \(\rho ^2\sum _iw_if_i^2+(1-\rho ^2)(\sum _iw_if_i)^2\le (\sum _iw_i|f_i|^p)^{2/p}\). Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\).
Let \(I\) be finite, let real weights satisfy \(\sum _iw_i=1\), and let \(\rho \in \mathbb {R}\). For real vectors \(f,g\), write \(m_h=\sum _iw_ih_i\) and \((T_\rho h)_i=\rho h_i+(1-\rho )m_h\). Then \(\sum _iw_i(T_\rho g)_i^2=\rho ^2\sum _iw_ig_i^2+(1-\rho ^2)m_g^2\) and \(\sum _iw_i(T_\rho f)_ig_i=\sum _iw_if_i(T_\rho g)_i\).
Let \(I\) be finite, \(w_i\ge 0\), \(\sum _iw_i=1\), \(q\gt 2\), and \(\rho \in \mathbb {R}\). Put \(p=q/(q-1)\) and \((T_\rho h)_i=\rho h_i+(1-\rho )\sum _jw_jh_j\). If \(\rho ^2\sum _iw_ig_i^2+(1-\rho ^2)(\sum _iw_ig_i)^2\le (\sum _iw_i|g_i|^p)^{2/p}\) for every \(g\), then \((\sum _iw_i|(T_\rho f)_i|^q)^{2/q}\le \sum _iw_if_i^2\) for every \(f\).
Let \(I\) be finite, \(w_i\gt 0\), \(\sum _iw_i=1\), \(1\lt p\lt 2\), and \(0\le \rho \le 1\). Define \(G(f)=\sum _iw_i|f_i|^p\) and \(F(f)=\rho ^2\sum _iw_if_i^2+(1-\rho ^2)(\sum _iw_if_i)^2\). There exists a nonnegative vector \(g\) with \(G(g)=1\) and \(F(f)\le F(g)\) whenever \(G(f)=1\).
Let \(I\) be finite, let \(\varphi :\mathbb {R}\to \mathbb {R}\) be strictly concave on \([0,\infty )\), and let \(f:I\to [0,\infty )\) satisfy \(\varphi (f_i)=c\) for every \(i\) and some \(c\in \mathbb {R}\). There exist \(a,b\ge 0\) such that every \(f_i\) equals \(a\) or \(b\).
Let \(I\) be finite, \(w:I\to \mathbb {R}\), and \(p,\rho ,t\in \mathbb {R}\). Define \(G(f)=\sum _iw_i|f_i|^p\) and \(F(f)=\rho ^2\sum _iw_if_i^2+(1-\rho ^2)(\sum _iw_if_i)^2\). Then \(G(tf)=|t|^pG(f)\) and \(F(tf)=t^2F(f)\) for every real vector \(f\).
Let \(I\) be finite, \(w:I\to \mathbb {R}\), \(p\ne 0\), and \(\rho \in \mathbb {R}\). Define \(G(f)=\sum _iw_i|f_i|^p\) and \(F(f)=\rho ^2\sum _iw_if_i^2+(1-\rho ^2)(\sum _iw_if_i)^2\). If \(G(f)\gt 0\) and \(c=G(f)^{-1/p}\), then \(G(cf)=1\) and \(F(cf)=F(f)/G(f)^{2/p}\).
Let \(I\) be finite, \(w:I\to \mathbb {R}\), \(p\gt 1\), \(\rho \in \mathbb {R}\), \(g_i\ge 0\), and \(h\in \mathbb {R}^I\). Define \(G(f)=\sum _iw_i|f_i|^p\) and \(F(f)=\rho ^2\sum _iw_if_i^2+(1-\rho ^2)(\sum _iw_if_i)^2\). At \(t=0\), the derivative of \(G(g+th)\) is \(p\sum _iw_ig_i^{p-1}h_i\), and that of \(F(g+th)\) is \(2[\rho ^2\sum _iw_ig_ih_i+(1-\rho ^2)(\sum _iw_ig_i)(\sum _iw_ih_i)]\).
Let \(I\) be finite, \(w_i\gt 0\), \(\sum _iw_i=1\), \(1\lt p\lt 2\), and \(0\le \rho \lt 1\). Define \(G(f)=\sum _iw_i|f_i|^p\) and \(F(f)=\rho ^2\sum _iw_if_i^2+(1-\rho ^2)(\sum _iw_if_i)^2\). If \(g_i\ge 0\), \(G(g)=1\), and \(F(f)\le F(g)\) whenever \(G(f)=1\), then \(F(g)\ge 1\) and \(F(g)g_i^{p-1}-\rho ^2g_i=(1-\rho ^2)\sum _jw_jg_j\) for every \(i\).
Let \(I\) be finite, \(w_i\gt 0\), \(\sum _iw_i=1\), \(1\lt p\lt 2\), and \(0\le \rho \lt 1\). Define \(G(f)=\sum _iw_i|f_i|^p\) and \(F(f)=\rho ^2\sum _iw_if_i^2+(1-\rho ^2)(\sum _iw_if_i)^2\). If \(g_i\ge 0\), \(G(g)=1\), and \(F(f)\le F(g)\) whenever \(G(f)=1\), then some \(a,b\ge 0\) satisfy \(g_i\in \{ a,b\} \) for every \(i\).
Let \(1\lt p\lt 2\), \(0\lt \alpha \lt \beta \), and \(\alpha ^p+\beta ^p=1\). Then \((\alpha ^p\beta ^{2-p}-\alpha ^{2-p}\beta ^p)/(\alpha ^2-\beta ^2)=\rho _*(p/(p-1),\alpha ^p)^2\). Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\).
Let \(1\lt p\lt 2\), \(0\lt \alpha \lt \beta \), and \(\alpha ^p+\beta ^p=1\). Put \(t=(\beta -\alpha )/(\beta +\alpha )\) and \(S=(1+t)^p+(1-t)^p\). Then \(S^{2/p}=4/(\alpha +\beta )^2\) and \(S^{2/p-1}((1+t)^{p-1}-(1-t)^{p-1})/t=2(\beta ^{p-1}-\alpha ^{p-1})/(\beta -\alpha )\).
Let \(1\lt p\lt 2\), \(0\lt \alpha \lt \beta \), \(\alpha ^p+\beta ^p=1\), and \(-1\le y\le 1\). Put \(r=(\alpha ^p\beta ^{2-p}-\alpha ^{2-p}\beta ^p)/(\alpha ^2-\beta ^2)\). Setting \(u=(1+y)/\alpha \) and \(v=(1-y)/\beta \), one has \((\alpha ^pu+\beta ^pv)^2+r\alpha ^p\beta ^p(u-v)^2\le ((1+y)^p+(1-y)^p)^{2/p}\).
Let \(1\lt p\lt 2\), \(0\lt \alpha \lt \beta \), \(\alpha ^p+\beta ^p=1\), and \(u,v\ge 0\). Put \(r=(\alpha ^p\beta ^{2-p}-\alpha ^{2-p}\beta ^p)/(\alpha ^2-\beta ^2)\). Then \((\alpha ^pu+\beta ^pv)^2+r\alpha ^p\beta ^p(u-v)^2\le (\alpha ^pu^p+\beta ^pv^p)^{2/p}\).
Let \(1\lt p\lt 2\), \(0\lt \alpha \lt \beta \), \(\alpha ^p+\beta ^p=1\), and \(u,v\in \mathbb {R}\). Put \(r=(\alpha ^p\beta ^{2-p}-\alpha ^{2-p}\beta ^p)/(\alpha ^2-\beta ^2)\). Then \((\alpha ^pu+\beta ^pv)^2+r\alpha ^p\beta ^p(u-v)^2\le (\alpha ^p|u|^p+\beta ^p|v|^p)^{2/p}\).
Let \(1\lt p\lt 2\), \(0\lt \lambda \lt 1/2\), and \(u,v\in \mathbb {R}\). Setting \(\rho =\rho _*(p/(p-1),\lambda )\), one has \((\lambda u+(1-\lambda )v)^2+\rho ^2\lambda (1-\lambda )(u-v)^2\le (\lambda |u|^p+(1-\lambda )|v|^p)^{2/p}\). Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\).
Let \(1\lt p\lt 2\), \(0\lt \alpha \lt \beta \), and \(\alpha ^p+\beta ^p=1\). Put \(r=(\alpha ^p\beta ^{2-p}-\alpha ^{2-p}\beta ^p)/(\alpha ^2-\beta ^2)\). For \(u=\beta /\alpha \) and \(v=\alpha /\beta \), one has \(u\ne v\), \(\alpha ^pu^p+\beta ^pv^p=1\), and \((\alpha ^pu+\beta ^pv)^2+r\alpha ^p\beta ^p(u-v)^2=1\).
Let \(1\lt p\lt 2\), \(0\lt \lambda \lt 1/2\), and \(\rho \ge 0\). If \((\lambda u+(1-\lambda )v)^2+\rho ^2\lambda (1-\lambda )(u-v)^2\le (\lambda |u|^p+(1-\lambda )|v|^p)^{2/p}\) for every \(u,v\in \mathbb {R}\), then \(\rho \le \rho _*(p/(p-1),\lambda )\). Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\).
Let \(q\gt 2\), \(0\lt \lambda \lt 1/2\), and \(u,v\in \mathbb {R}\). Put \(\rho =\rho _*(q,\lambda )\) and \(m=\lambda u+(1-\lambda )v\). Then \((\lambda |\rho u+(1-\rho )m|^q+(1-\lambda )|\rho v+(1-\rho )m|^q)^{2/q}\le \lambda u^2+(1-\lambda )v^2\). Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\).
Let \(q\gt 2\), \(0\lt \lambda \lt 1/2\), \(\lambda \le \gamma \le 1-\lambda \), and \(u,v\in \mathbb {R}\). Put \(p=q/(q-1)\) and \(\rho =\rho _*(q,\lambda )\). Then \((\gamma u+(1-\gamma )v)^2+\rho ^2\gamma (1-\gamma )(u-v)^2\le (\gamma |u|^p+(1-\gamma )|v|^p)^{2/p}\). Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\).
Let \(q\gt 2\), \(0\lt \lambda \lt 1/2\), and let \(X\) be measurable on a probability space with \(\Pr [X=1-\lambda ]=\lambda \) and \(\Pr [X=-\lambda ]=1-\lambda \). Then \(H_{q/(q-1),2,\rho _*(q,\lambda )}(X)\) holds. Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(q\gt 2\), \(0\lt \lambda \lt 1/2\), and let \(X\) be measurable on a probability space with \(\Pr [X=1-\lambda ]=\lambda \) and \(\Pr [X=-\lambda ]=1-\lambda \). If \(H_{q/(q-1),2,\rho }(X)\) holds, then \(\rho \le \rho _*(q,\lambda )\). Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(q\gt 2\), \(0\lt \lambda \lt 1/2\), and let \(X\) be measurable on a probability space with \(\Pr [X=1-\lambda ]=\lambda \) and \(\Pr [X=-\lambda ]=1-\lambda \). Then \(H_{2,q,\rho _*(q,\lambda )}(X)\) holds. Here \(\rho _*(q,\lambda )=\sqrt{\sinh (\ell /q)/\sinh (\ell /(q/(q-1)))}\) with \(\ell =\log ((1-\lambda )/\lambda )\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(X\in L^4\) have a symmetric law on a probability space, \(\| X\| _2=1\), and \(\| X\| _4=C\gt 0\). For \(0\le \rho \lt 1\), \(H_{2,4,\rho }(X)\) holds if and only if \(\rho \le \min (1/\sqrt3,1/C)\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \(q\gt 2\) and let \(X\in L^q\) have a symmetric law on a probability space, with \(\| X\| _2=1\) and \(\| X\| _q=C\gt 0\). Then \(H_{2,q,\rho }(X)\) holds for \(\rho =1/(C\sqrt{q-1})\). Write \(H_{p,q,\rho }(X)\) for \(1\le p\le q\le \infty \), \(0\le \rho \lt 1\), \(X\in L^q(\mu )\), and \(\| a+\rho bX\| _q\le \| a+bX\| _p\) for all \(a,b\in \mathbb {R}\).
Let \((A,\mu )\) and \((B,\nu )\) be probability spaces, let \(f:A\times B\to [0,\infty ]\) be measurable, and let \(r\ge 1\). Then \((\int _B(\int _A f(x,y)\, d\mu (x))^r\, d\nu (y))^{1/r}\le \int _A(\int _B f(x,y)^r\, d\nu (y))^{1/r}\, d\mu (x)\), with nonnegative extended integrals.
Let \((A,\mu )\) and \((B,\nu )\) be probability spaces, let \(g:A\times B\to [0,\infty ]\) be measurable, and let \(0\lt p\le q\lt \infty \). Then \((\int _B(\int _A g^p\, d\mu )^{q/p}\, d\nu )^{1/q}\le (\int _A(\int _B g^q\, d\nu )^{p/q}\, d\mu )^{1/p}\).
Let \(\mu ,\nu \) be probability measures on \(\mathbb {R}\), \(1\le p\le q\lt \infty \), and \(\rho \in \mathbb {R}\). Suppose \(\| a+\rho bx\| _{L^q(\theta )}\le \| a+bx\| _{L^p(\theta )}\) for every \(a,b\in \mathbb {R}\) and each \(\theta \in \{ \mu ,\nu \} \). Then \(\| a+\rho b(x+y)\| _{L^q(\mu \times \nu )}\le \| a+b(x+y)\| _{L^p(\mu \times \nu )}\) for every \(a,b\).
Let \(f\) be an almost everywhere strongly measurable real function on an arbitrary measure space \((A,\mu )\), let \(C\in [0,\infty ]\) and \(N\in \mathbb N\). If \(\| f\| _n\le C\) for every integer \(n\ge N\), then \(\| f\| _\infty \le C\).
Let \(\mu ,\nu \) be probability measures on \(\mathbb {R}\), \(1\le p\le q\le \infty \), and \(\rho \in \mathbb {R}\). Suppose \(\| a+\rho bx\| _{L^q(\theta )}\le \| a+bx\| _{L^p(\theta )}\) for every \(a,b\in \mathbb {R}\) and each \(\theta \in \{ \mu ,\nu \} \). Then \(\| a+\rho b(x+y)\| _{L^q(\mu \times \nu )}\le \| a+b(x+y)\| _{L^p(\mu \times \nu )}\) for every \(a,b\), with the essential supremum at infinity.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(J,S\subseteq [n]\), \(f:\Omega \to \mathbb R\), and \(x\in \Omega \). Then \(E_J(f_S)(x)=f_S(x)\) if \(S\subseteq J\), and \(E_J(f_S)(x)=0\) otherwise.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S,T\subseteq [n]\), \(f:\Omega \to \mathbb R\), and \(x\in \Omega \). The \(S\)-component of \(f_T\) equals \(f_S(x)\) when \(S=T\), and equals zero otherwise.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S\subseteq [n]\), \(\rho \in \mathbb R\), and \(f:\Omega \to \mathbb R\). The \(S\)-component of \(T_\rho f\) is \(\rho ^{|S|}f_S\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) and \(f:\Omega \to \mathbb R\). Its orthogonal components satisfy \(\sum _{S\subseteq [n]}f_S(x)=f(x)\) for every \(x\in \Omega \).
Let \(s\) be finite with weights \(w_i\ge 0\) summing to \(1\), values \(F_i\in \mathbb R\), \(q\gt 0\), and \(0\le c\le 1\). If \(|1-cz|^q\le |1+z|^q-q(1+c)z\) for every real \(z\), then with \(a=\sum _{i\in s}w_iF_i\), \(\sum _{i\in s}w_i|a-c(F_i-a)|^q\le \sum _{i\in s}w_i|F_i|^q\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(J\subseteq [n]\), \(f:\Omega \to \mathbb R\), \(q\gt 0\), and \(0\le c\le 1\). If \(|1-cz|^q\le |1+z|^q-q(1+c)z\) for every real \(z\), then \(\mathbb E_P|E_Jf-c(f-E_Jf)|^q\le \mathbb E_P|f|^q\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(J\subseteq [n]\), \(f:\Omega \to \mathbb R\), and \(q\ge 1\). Then \(\mathbb E_P|E_Jf|^q\le \mathbb E_P|f|^q\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(k\in \mathbb N\), \(f:\Omega \to \mathbb R\) of degree at most \(k\), and \(c\in \mathbb R\). Then \(x\mapsto f(x)-c\) has degree at most \(k\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\) of degree at most \(k\in \mathbb N\), and \(\operatorname {Var}_P(f)\gt 0\). Then \(k\gt 0\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) and \(f:\Omega \to \mathbb R\) with orthogonal components \(f_S\). Then \(\operatorname {Var}_P(f)=\sum _{\varnothing \ne S\subseteq [n]}\mathbb E_Pf_S^2\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(J,S\subseteq [n]\), and \(f:\Omega \to \mathbb R\). The \(S\)-component of \(E_Jf\) equals \(f_S\) when \(S\subseteq J\), and zero otherwise.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), a coordinate \(i\), and \(S\subseteq [n]\). The \(S\)-component of \(L_if\) is \(f_S\) if \(i\in S\), and zero otherwise.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(J\subseteq [n]\), and \(f:\Omega \to \mathbb R\). Then \(\mathbb E_P(f-E_Jf)^2=\sum _{S\nsubseteq J}\mathbb E_Pf_S^2\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), and \(k\in \mathbb N\). Then \((k+1)\sum _{|S|\gt k}\mathbb E_Pf_S^2\le I(f)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \{ -1,1\} \), \(J\subseteq [n]\), \(k\in \mathbb N\), and \(a,b\in \mathbb R\). If \(\sum _{|S|\gt k}\mathbb E_Pf_S^2\le a\) and \(\sum _{|S|\le k,\ S\nsubseteq J}\mathbb E_Pf_S^2\le b\), then there exists \(g:\Omega \to \{ -1,1\} \) depending only on \(J\) with \(\Pr _P(f\ne g)\le a+b\).
Let \((\Omega ,\mu )\) be a probability space, \(q\ge 2\), and \(0\le c\le 1\) satisfy \(|1-cz|^q\le |1+z|^q-q(1+c)z\) for every real \(z\). If \(X\in L^q(\mu )\) is real with \(\int X\, d\mu =0\), then for every \(a\in \mathbb R\), \(\| a-cX\| _{L^q(\mu )}\le \| a+X\| _{L^q(\mu )}\).
For every \(q\ge 2\) there exists \(0\lt c\le 1\), with \(c=2/5\) if \(q=4\), such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), every coordinate \(i\), and every real function \(f\) on its outcome space, \(\| N_{i,-c}f\| _{q,P}\le \| f\| _{q,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(N_{i,a}h=E_{[n]\setminus \{ i\} }h+a(h-E_{[n]\setminus \{ i\} }h)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), a coordinate \(i\), \(\rho \in \mathbb R\), and \(f,g:\Omega \to \mathbb R\). Then \(\mathbb E_P[(N_{i,\rho }f)g]=\mathbb E_P[fN_{i,\rho }g]\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(N_{i,a}h=E_{[n]\setminus \{ i\} }h+a(h-E_{[n]\setminus \{ i\} }h)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), a coordinate \(i\), \(0\le \rho \le 1\), \(q\ge 1\), and \(f:\Omega \to \mathbb R\). Then \(\| N_{i,\rho }f\| _{q,P}\le \| f\| _{q,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(N_{i,a}h=E_{[n]\setminus \{ i\} }h+a(h-E_{[n]\setminus \{ i\} }h)\).
For every \(q\gt 1\) there exists \(0\lt c\le 1\), with \(c=2/5\) when \(q\in \{ 4,4/3\} \), such that every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) satisfies \(\| N_{i,-c}f\| _{q,P}\le \| f\| _{q,P}\) for every coordinate \(i\) and real function \(f\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(N_{i,a}h=E_{[n]\setminus \{ i\} }h+a(h-E_{[n]\setminus \{ i\} }h)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(J\subseteq [n]\), rates \(r_i\in \mathbb R\), and \(f:\Omega \to \mathbb R\). Define \(T_{J,\boldsymbol r}f(x)=\sum _S(\prod _{j\in S\cap J}r_j)f_S(x)\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(J\subseteq [n]\), rates \(r_i\in \mathbb R\), \(i\notin J\), and \(f:\Omega \to \mathbb R\). For \(T_{J,\boldsymbol r}f=\sum _S(\prod _{j\in S\cap J}r_j)f_S\), \(T_{J\cup \{ i\} ,\boldsymbol r}f=N_{i,r_i}(T_{J,\boldsymbol r}f)\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(N_{i,a}h=E_{[n]\setminus \{ i\} }h+a(h-E_{[n]\setminus \{ i\} }h)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(q\in \mathbb R\), and rates \(r_i\in \mathbb R\). If \(\| N_{i,r_i}g\| _{q,P}\le \| g\| _{q,P}\) for every coordinate \(i\) and every real \(g:\Omega \to \mathbb R\), then \(\| T_{\boldsymbol r}f\| _{q,P}\le \| f\| _{q,P}\) for every real \(f\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_{\boldsymbol s}h=\sum _{S\subseteq [n]}(\prod _{i\in S}s_i)h_S\) for real coordinate rates \(s_i\); \(N_{i,a}h=E_{[n]\setminus \{ i\} }h+a(h-E_{[n]\setminus \{ i\} }h)\); \(\| h\| _{q,P}:=(\mathbb E_P|h|^q)^{1/q}\) as a totalized moment expression for every real \(q\), with \(1/0=0\), \(0^0=1\), and \(0^b=0\) for \(b\ne 0\).
For every \(q\gt 1\) there exists \(0\lt c\le 1\), with \(c=2/5\) when \(q\in \{ 4,4/3\} \), such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), every sign vector \(s\in \{ -1,1\} ^n\), and every real \(f\) on its outcome space, \(\| T_{c\boldsymbol s}f\| _{q,P}\le \| f\| _{q,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(T_{\boldsymbol s}h=\sum _{S\subseteq [n]}(\prod _{i\in S}s_i)h_S\) for real coordinate rates \(s_i\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), \(r\in \{ 0,1\} ^n\), and \(x\in \Omega \). Define \(R_rf(x)=\sum _{S\subseteq [n]}\chi _S(r)f_S(x)\), where \(f_S\) is the orthogonal component and \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(q\in \mathbb R\), and \(f:\Omega \to \mathbb R\). For \(R_rf(x)=\sum _S\chi _S(r)f_S(x)\) and uniform independent \(r\in \{ 0,1\} ^n\), define \(\| Rf\| _q=(\mathbb E_{x\sim P}\mathbb E_r|R_rf(x)|^q)^{1/q}\) using real powers.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), a coordinate \(i\), \(\rho \in \mathbb R\), and \(f:\Omega \to \mathbb R\). Define \(N_{i,\rho }f=f-L_if+\rho L_if\), where \(L_i\) subtracts the conditional expectation over coordinate \(i\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), rates \(r_i\in \mathbb R\), and \(f:\Omega \to \mathbb R\). Define \(T_{\boldsymbol r}f(x)=\sum _S(\prod _{i\in S}r_i)f_S(x)\) using the orthogonal components \(f_S\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), a coordinate \(i\), and \(x\in \Omega \). Define \(\mathcal I_i(f,x)=\sum _{S\ni i}f_S(x)^2\) using the orthogonal components of \(f\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), \(\tau \in \mathbb R\), and \(x\in \Omega \). Define \(J_\tau (f,x)=\{ i\in [n]:\tau \le \sum _{S\ni i}f_S(x)^2\} \).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), and \(x\in \Omega \). If \(R_rf(x)=\sum _S\chi _S(r)f_S(x)\) and \(r\) is uniform on the Boolean cube, then \(\mathbb E_r(R_rf(x))^2=\sum _Sf_S(x)^2\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\) for \(r\in \{ 0,1\} ^n\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) and \(f:\Omega \to \mathbb R\). Then \((\mathbb E_{x\sim P}\mathbb E_r|\sum _S\chi _S(r)f_S(x)|^2)^{1/2}=\| f\| _{2,P}\), where \(r\) is uniform on the Boolean cube. For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\) for \(r\in \{ 0,1\} ^n\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube, \(k\in \mathbb N\), and \(q\gt 1\). The Fourier projection \(f^{\le k}\) satisfies \(\| f^{\le k}\| _q\le (\sqrt{q-1})^k\| f\| _q\) when \(q\ge 2\), and \(\| f^{\le k}\| _q\le (1/\sqrt{q-1})^k\| f\| _q\) when \(q\le 2\). For a real function \(h\) on the cube, use \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\), \(\widehat h(S)=\mathbb E_x[h(x)\chi _S(x)]\), \(h^{\le k}=\sum _{|S|\le k}\widehat h(S)\chi _S\), \(\| h\| _p=(\mathbb E_x|h(x)|^p)^{1/p}\), with uniform expectation \(\mathbb E_x\) and natural powers \(0^0=1\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube, \(k\in \mathbb N\), and \(0\lt \rho \le 1\). Then \(\| f^{\le k}\| _4\le (\sqrt3/\rho )^k\| T_\rho f\| _4\). For a real function \(h\) on the cube, use \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\), \(\widehat h(S)=\mathbb E_x[h(x)\chi _S(x)]\), \(h^{\le k}=\sum _{|S|\le k}\widehat h(S)\chi _S\), \(\| h\| _p=(\mathbb E_x|h(x)|^p)^{1/p}\), \(T_a h=\sum _{S\subseteq [n]}a^{|S|}\widehat h(S)\chi _S\), with uniform expectation \(\mathbb E_x\) and natural powers \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), and \(r\in \{ 0,1\} ^n\). For \(s_i=(-1)^{r_i}\), \(R_rf(x)=T_{\boldsymbol s}f(x)\) for every \(x\in \Omega \). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\) for \(r\in \{ 0,1\} ^n\); \(R_rh=\sum _{S\subseteq [n]}\chi _S(r)h_S\); \(T_{\boldsymbol s}h=\sum _{S\subseteq [n]}(\prod _{i\in S}s_i)h_S\) for real coordinate rates \(s_i\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), and \(x\in \Omega \). For uniform \(r\in \{ 0,1\} ^n\), \(R_rf=\sum _S\chi _S(r)f_S\), and \(M(r)=\{ i:r_i=0\} \), one has \(\mathbb E_r[E_{M(r)}(R_rf)(x)]=T_{1/2}f(x)\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\) for \(r\in \{ 0,1\} ^n\); \(R_rh=\sum _{S\subseteq [n]}\chi _S(r)h_S\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), and \(q\ge 1\). Then \(\| T_{1/2}f\| _{q,P}\le (\mathbb E_{x\sim P}\mathbb E_r|\sum _S\chi _S(r)f_S(x)|^q)^{1/q}\), where \(r\) is uniform on the Boolean cube. For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\) for \(r\in \{ 0,1\} ^n\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
For every \(q\ge 2\) there exists \(0\lt c\le 1\), with \(c=2/5\) if \(q=4\), such that every probability space \((\Omega ,\mu )\), every real \(X\in L^q(\mu )\) of mean zero, and every real \(a\) satisfy \(\| a-cX\| _q\le \| a+X\| _q\). Here \(\| Y\| _q=(\int _\Omega |Y|^q\, d\mu )^{1/q}\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(c\in \mathbb R\), \(f:\Omega \to \mathbb R\), and \(r\in \{ 0,1\} ^n\). Then \(R_r(T_cf)=T_{\boldsymbol s}f\) for \(s_i=c(-1)^{r_i}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\) for \(r\in \{ 0,1\} ^n\); \(R_rh=\sum _{S\subseteq [n]}\chi _S(r)h_S\); \(T_{\boldsymbol s}h=\sum _{S\subseteq [n]}(\prod _{i\in S}s_i)h_S\) for real coordinate rates \(s_i\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f,g:\Omega \to \mathbb R\), and \(q\gt 0\). If \(\| R_rf\| _{q,P}\le \| g\| _{q,P}\) for every sign vector \(r\), then \((\mathbb E_{x\sim P}\mathbb E_r|R_rf(x)|^q)^{1/q}\le \| g\| _{q,P}\), with uniform signs \(r\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\) for \(r\in \{ 0,1\} ^n\); \(R_rh=\sum _{S\subseteq [n]}\chi _S(r)h_S\).
For every \(q\gt 1\) there exists \(0\lt c\le 1\), with \(c=2/5\) if \(q\in \{ 4,4/3\} \), such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\) and every real \(f\) on its outcome space, \((\mathbb E_{x\sim P}\mathbb E_r|R_r(T_cf)(x)|^q)^{1/q}\le \| f\| _{q,P}\), where \(r\) is uniform on the Boolean cube. For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(\chi _S(r)=\prod _{i\in S}(-1)^{r_i}\) for \(r\in \{ 0,1\} ^n\); \(R_rh=\sum _{S\subseteq [n]}\chi _S(r)h_S\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(\rho \ne 0\), and \(f:\Omega \to \mathbb R\). Then \(T_\rho (T_{1/\rho }f)=f\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(\rho \in \mathbb R\), and \(f:\Omega \to \mathbb R\). Anisotropic noise with every coordinate rate equal to \(\rho \) satisfies \(T_{(\rho ,\ldots ,\rho )}f(x)=T_\rho f(x)\) for every \(x\in \Omega \). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(T_a h=\sum _{S\subseteq [n]}a^{|S|}h_S\) for every real rate \(a\); \(T_{\boldsymbol s}h=\sum _{S\subseteq [n]}(\prod _{i\in S}s_i)h_S\) for real coordinate rates \(s_i\). The powers indexed by subset cardinalities are natural powers, including \(0^0=1\).
For every \(q\ge 2\) there exists \(C\ge 1\), with \(C=5\sqrt3\) if \(q=4\), such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), every \(k\in \mathbb N\), and every real \(f\) on its outcome space, \(\| f^{\le k}\| _{q,P}\le C^k\| f\| _{q,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(h^{\le k}=\sum _{|S|\le k}h_S\).
For every \(q\gt 1\) there exists \(C\ge 1\), with \(C=5\sqrt3\) when \(q\in \{ 4,4/3\} \), such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), every \(k\in \mathbb N\), and every real \(f\) on its outcome space, \(\| f^{\le k}\| _{q,P}\le C^k\| f\| _{q,P}\). For a real function \(h\) on \(\Omega \), use \(E_Uh(x)=\mathbb E_P[h(X)\mid X_U=x_U]\) for \(X\sim P\) and \(U\subseteq [n]\); \(h_S=\sum _{U\subseteq S}(-1)^{|S|-|U|}E_Uh\) for \(S\subseteq [n]\) (the Hoeffding–Efron–Stein components); \(\| h\| _{p,P}=(\mathbb E_P|h|^p)^{1/p}\) for \(p\gt 0\); \(h^{\le k}=\sum _{|S|\le k}h_S\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), \(x\in \Omega \), and \(S\subseteq [n]\). The Walsh coefficient at \(S\) of \(r\mapsto R_rf(x)=\sum _T\chi _T(r)f_T(x)\) is \(f_S(x)\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f,g:\Omega \to \mathbb R\), \(q\gt 0\), and \(A\ge 0\). If \(\| r\mapsto R_rf(x)\| _q\le A\| r\mapsto R_rg(x)\| _q\) on the uniform sign cube for every \(x\), then \(\| Rf\| _q\le A\| Rg\| _q\), where \(\| Rh\| _q=(\mathbb E_{x\sim P}\mathbb E_r|R_rh(x)|^q)^{1/q}\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(S\subseteq [n]\), \(k\in \mathbb N\), and \(f:\Omega \to \mathbb R\). The \(S\)-component of \(f^{\le k}\) equals \(f_S\) if \(|S|\le k\), and zero otherwise.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(\rho \in \mathbb R\), \(f:\Omega \to \mathbb R\), and \(x\in \Omega \). Then \(r\mapsto R_r(T_\rho f)(x)\) equals \(T_\rho (r\mapsto R_rf(x))\), with the noise on the right acting on the uniform Boolean sign cube.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(k\in \mathbb N\), \(f:\Omega \to \mathbb R\), and \(x\in \Omega \). Then \(r\mapsto R_r(f^{\le k})(x)\) is the Fourier projection to levels at most \(k\) of the function \(r\mapsto R_rf(x)\) on the uniform Boolean sign cube.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(k\in \mathbb N\), and \(f,g:\Omega \to \mathbb R\). Then \(\mathbb E_P[f^{\le k}g]=\mathbb E_P[fg^{\le k}]\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(\rho \in \mathbb R\), \(k\in \mathbb N\), and \(f:\Omega \to \mathbb R\). Then \(T_\rho (f^{\le k})=(T_\rho f)^{\le k}\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(q\gt 1\), \(A\gt 0\), and an operator \(T\) on real functions on \(\Omega \). If \(\mathbb E_P[(Tf)g]=\mathbb E_P[fTg]\) for every \(f,g\) and \(\| Tg\| _{q/(q-1),P}\le A\| g\| _{q/(q-1),P}\) for every \(g\), then \(\| Tf\| _{q,P}\le A\| f\| _{q,P}\) for every \(f\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), \(T\subseteq [n]\), \(x\in \Omega \), and \(\tau \in \mathbb R\). The restriction \((T,x)\) is a \(\tau \)-booster if either \(\tau \gt 0\) and \(\mathbb E_Pf+\tau \le E_Tf(x)\), or \(\tau \lt 0\) and \(E_Tf(x)\le \mathbb E_Pf+\tau \).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), sets \(J_j\subseteq [n]\) and tests \(b_j:\Omega \to \{ 0,1\} \) indexed by \(j\in \{ 0,\ldots ,m-1\} \). Define \(J(x)=\bigcup _{j:b_j(x)=1}J_j\).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), sets \(J_j\subseteq [n]\) and tests \(b_j:\Omega \to \{ 0,1\} \) for \(j\lt m\). The observation at \(x\) records \(x_i\) when \(i\in \bigcup _{j:b_j(x)=1}J_j\), and records a distinguished missing symbol otherwise.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(h:\Omega \to \mathbb R\), and \(K\in \mathbb R\). Say that \(h\) is a \(K\)-pseudo-junta if there exist finitely many sets \(J_j\subseteq [n]\), tests \(b_j:\Omega \to \{ 0,1\} \) depending only on \(J_j\), and a real function of the coordinate observations such that \(h(x)\) is determined by the values on \(J(x)=\bigcup _{j:b_j(x)=1}J_j\), including knowledge of \(J(x)\) itself, and \(\mathbb E_P|J(x)|\le K\).
For every \(p\in \mathbb R\) and an event \(A\subseteq \{ 0,1\} ^n\), define \(\Pr _p(A)=\sum _x\prod _i[p\mathbf1_{x_i=1}+(1-p)\mathbf1_{x_i=0}]\mathbf1_A(x)\). For \(0\le p\le 1\) this is independent Bernoulli-\(p\) probability; the formula defines a polynomial for every real \(p\).
Let \(e:[n]\to \binom {[v]}2\) be a bijection and \(f:\{ 0,1\} ^n\to \{ 0,1\} \). Say that \(f\) is a graph property if for every vertex permutation \(\sigma \) and inputs \(x,y\) satisfying \(y_j=x_i\) whenever \(e(j)=\sigma (e(i))\), one has \(f(y)=f(x)\).
There exists an absolute \(B\gt 0\) such that for every \(n\ge 2\), nonconstant increasing transitively symmetric \(f:\{ 0,1\} ^n\to \{ 0,1\} \), \(0\lt p_c\le 1/2\) with \(\Pr _{p_c}(f=1)=1/2\), and \(0\lt \varepsilon \lt 1/4\), the number \(\eta =B\log (1/\varepsilon )\log (1/p_c)/\log n\) satisfies: if \(\eta \le 1/2\), then \(\Pr _{p_c(1-\eta )}(f=1)\le \varepsilon \) and \(\Pr _{p_c(1+\eta )}(f=1)\ge 1-\varepsilon \).
There exists \(w:\mathbb R\times \mathbb R\to \mathbb N\) such that whenever the \(n\) variables bijectively index the edges of a labeled simple graph on \(v\) vertices, \(f:\{ 0,1\} ^n\to \{ 0,1\} \) is an increasing graph property, \(0\lt p\le 1/2\), \(K\gt 0\), \(0\lt \varepsilon \lt 1\), and \(I_p(f)\le K\), there is a monotone DNF \(d\) of width at most \(w(K,\varepsilon )\) with \(\Pr _p(f\ne d)\le \varepsilon \). Here \(I_p(f)=4p(1-p)\) times the sum of pivotal probabilities.
There are absolute \(C,D\gt 0\) such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), every \(f:\Omega \to \{ -1,1\} \), and every real \(K\) satisfying \(I(f)\le K\) and \(\operatorname {Var}_P(f)\ge 1/100\), there exists \(\tau \in \mathbb R\) with \(e^{-CK^2}\le |\tau |\le \Pr _P(\exists T\subseteq [n],\ |T|\le DK:\ (T,x)\text{ is a }\tau \text{-booster})\). A \(\tau \)-booster has \(E_Tf(x)\ge \mathbb E_Pf+\tau \) for \(\tau \gt 0\), or \(E_Tf(x)\le \mathbb E_Pf+\tau \) for \(\tau \lt 0\).
There is an absolute \(C\gt 0\) such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), every \(f:\Omega \to \{ -1,1\} \), \(K\ge 0\) with \(I(f)\le K\), and \(\varepsilon \gt 0\), there is \(h:\Omega \to \{ -1,1\} \) with \(\Pr _P(f\ne h)\le \varepsilon \) that is determined by locally triggered coordinate observations exposing at most \(\exp (CK^3/\varepsilon ^3)\) coordinates in expectation.
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(h:\Omega \to \{ -1,1\} \), and \(K\in \mathbb R\). If \(h\) is determined by locally triggered coordinate observations whose expected exposed size is at most \(K\), with each trigger depending only on its declared coordinate set, then \(I(h)\le 4K\).
There are absolute \(C,D\gt 0\) such that for every product \(P\) of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), every \(f:\Omega \to \{ -1,1\} \), and \(0\lt \varepsilon \lt 1/2\), putting \(k=I(f)/\varepsilon \), there exist sets \(J(x)\subseteq [n]\) satisfying \(|J(x)|\le e^{Ck}\) and \(|\{ S\subseteq J(x):|S|\le k\} |\le e^{Dk^2}\) for every \(x\), and \(\mathbb E_P\sum _{S\notin \{ T\subseteq J(x):|T|\le k\} }f_S(x)^2\le 2\varepsilon \).
Let \(P\) be a product of \(n\) finite probability spaces with strictly positive atom masses and outcome space \(\Omega =\prod _i\Omega _i\), \(f:\Omega \to \mathbb R\), \(x\in \Omega \), a coordinate \(i\), and \(\tau \gt 0\). If \(\sum _{S\ni i}f_S(x)^2\lt \tau \), put \(g=T_{2/5}L_if\) and \(h(r)=\sum _S\chi _S(r)g_S(x)\); then on the uniform sign cube \(\| T_{1/\sqrt3}h\| _2^2\le \tau ^{1/3}\| h\| _{4/3}^{4/3}\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. Write \(x^{\{ i\} }\) for \(x\) with bit \(i\) flipped, \(I_i(h)=\mathbb E[(h(x)-h(x^{\{ i\} }))^2]/4\), and \(I(h)=\sum _i I_i(h)\). For every real function \(h\) on the cube and coordinate \(i\in [n]\), \(I_i(|h|)\le I_i(h)\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Write \(x^{\{ i\} }\) for \(x\) with bit \(i\) flipped, \(I_i(h)=\mathbb E[(h(x)-h(x^{\{ i\} }))^2]/4\), and \(I(h)=\sum _i I_i(h)\). For every real cube function \(g\), \(2(\mathbb E[g^2]-(\mathbb Eg)^2)\le I(g)+\sum _{|S|=1}\widehat g(S)^2\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(g\) be a real cube function and \(a\in \mathbb R\). If \(|g(x)-g(\bar x)|\le 2|a|\) for every \(x\), where \(\bar x_i=1-x_i\), then \(\sum _{|S|=1}\widehat g(S)^2\le a^2\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). Suppose \(a_S=0\) whenever \(|S|\gt 1\). For every \(x\in \{ 0,1\} ^n\), \(P(x)+P(\bar x)=2a_\varnothing \), where \(\bar x_i=1-x_i\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). Write \(x^{\{ i\} }\) for \(x\) with bit \(i\) flipped, \(I_i(h)=\mathbb E[(h(x)-h(x^{\{ i\} }))^2]/4\), and \(I(h)=\sum _i I_i(h)\). If \(a_S=0\) for \(|S|\gt 1\), then \(I(P)+a_\varnothing ^2=\sum _S a_S^2\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). If \(a_S=0\) for \(|S|\gt 1\), then \(\sum _{S\subseteq [n]}a_S^2\le 2(\mathbb E|P|)^2\).
For a natural number \(n\), let \(I=\{ 0,\ldots ,n-1\} \) be the coordinate set. A multilinear polynomial on the \(n\)-dimensional Boolean cube is represented by an arbitrary coefficient family \(p:\mathcal P(I)\to \mathbb R\), assigning a real coefficient \(p(S)\) to each subset \(S\subseteq I\).
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and let \(p:\mathcal P(I)\to \mathbb R\) be a coefficient family. For \(x\in \{ 0,1\} ^I\), its polynomial value is
where the product over the empty set is \(1\).
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(p:\mathcal P(I)\to \mathbb R\). The threshold function associated with \(p\) assigns to each \(x\in \{ 0,1\} ^I\) the value \(1\) when \(\sum _{S\subseteq I}p(S)\prod _{i\in S}(-1)^{x_i}\ge 0\), and \(-1\) otherwise. In particular, a zero polynomial value is assigned \(1\).
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), \(f:\{ 0,1\} ^I\to \mathbb R\), and \(p:\mathcal P(I)\to \mathbb R\). The family \(p\) is a polynomial threshold representation of \(f\) precisely when, for every \(x\in \{ 0,1\} ^I\), \(f(x)=\sigma (\sum _{S\subseteq I}p(S)\prod _{i\in S}(-1)^{x_i})\), where \(\sigma (r)=1\) for \(r\ge 0\) and \(\sigma (r)=-1\) for \(r\lt 0\).
Let \(n,k\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). The function \(f\) is a polynomial threshold function of degree at most \(k\) precisely when there exists \(p:\mathcal P(I)\to \mathbb R\) such that \(p(S)=0\) whenever \(|S|\gt k\) and
where \(\sigma (r)=1\) for \(r\ge 0\) and \(-1\) for \(r\lt 0\).
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). The function \(f\) is a linear threshold function precisely when it has a polynomial threshold representation of degree at most one: there exists \(p:\mathcal P(I)\to \mathbb R\) with \(p(S)=0\) for \(|S|\gt 1\) and \(f(x)=\sigma (\sum _{S\subseteq I}p(S)\prod _{i\in S}(-1)^{x_i})\) for every \(x\), where \(\sigma (r)=1\) for \(r\ge 0\) and \(-1\) for \(r\lt 0\).
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), \(f:\{ 0,1\} ^I\to \mathbb R\), and \(a:I\to \mathbb R\). The weights \(a\) give a homogeneous linear threshold representation of \(f\) precisely when, for every \(x\in \{ 0,1\} ^I\), the sum \(L_a(x)=\sum _{i\in I}a_i(-1)^{x_i}\) is nonzero and \(f(x)=\sigma (L_a(x))\), where \(\sigma (r)=1\) for \(r\ge 0\) and \(-1\) for \(r\lt 0\).
Let \(n,s\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). The function \(f\) is a majority of \(s\) signed parities precisely when there are subsets \(S_0,\ldots ,S_{s-1}\subseteq I\) and signs \(\varepsilon _j\in \{ -1,1\} \) such that
where repetitions of the subsets are allowed and \(\sigma (r)=1\) for \(r\ge 0\), \(-1\) for \(r\lt 0\); thus ties, including the empty sum when \(s=0\), give the value \(1\).
Let \(n,k\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). Its Fourier weight on levels at most \(k\) is
Let \(n,k\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). Its Fourier weight on levels strictly above \(k\) is
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). Its spectral one-norm is
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). The degree-one part of \(f\) is the function
The constant Fourier coefficient is excluded.
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). The maximum influence is
where \(x^{\oplus i}\) is obtained by flipping coordinate \(i\), and the supremum over the empty coordinate set is assigned \(0\) when \(n=0\).
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), \(f:\{ 0,1\} ^I\to \mathbb R\), and \(\rho \in \mathbb R\). The noise stability at \(\rho \) is
where \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat f(S)=2^{-n}\sum _{x\in \{ 0,1\} ^I}f(x)\chi _S(x)\); this quantity is defined for every real \(\rho \).
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), \(f:\{ 0,1\} ^I\to \mathbb R\), and \(\delta \in \mathbb R\). The noise sensitivity is the real number
where \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat f(S)=2^{-n}\sum _{x\in \{ 0,1\} ^I}f(x)\chi _S(x)\); neither \(\delta \in [0,1]\) nor \(f(x)\in \{ -1,1\} \) is required.
For \(n\in \mathbb N\), a point of \(\{ 0,1\} ^{n+n}\) is written as two blocks \((u,v)\) with \(u,v\in \{ 0,1\} ^n\). The inner-product-mod-two function is \(\operatorname {IP}_n(u,v)=\prod _{i=0}^{n-1}(-1)^{u_i v_i}=(-1)^{\sum _{i=0}^{n-1}u_i v_i}\), so it is \(-1\) exactly when the number of coordinates where both bits are \(1\) is odd; the empty product for \(n=0\) is \(1\).
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), \(J\subseteq I\), and \(b\in \{ 0,1\} ^I\). The subcube indicator associated with \(J\) and \(b\) is the function taking value \(1\) at \(x\in \{ 0,1\} ^I\) if \(x_i=b_i\) for every \(i\in J\), and value \(0\) otherwise.
Let \(n\in \mathbb N\) and \(t\in \mathbb R\). The normalized Hamming-ball indicator takes value \(1\) at \(x\in \{ 0,1\} ^n\) precisely when \(t\le (\sum _{i=0}^{n-1}(-1)^{x_i})/\sqrt n\), and value \(0\) otherwise, including the boundary in the value-\(1\) region. When \(n=0\), the quotient is assigned \(0\), so the function takes value \(1\) exactly when \(t\le 0\).
Let \(n\in \mathbb N\) and \(g:\{ 0,1\} ^n\to \mathbb R\), and define \(\sigma (t)=1\) when \(t\ge 0\) and \(\sigma (t)=-1\) when \(t\lt 0\). Then the function \(x\mapsto \sigma (g(x))\) takes values in \(\{ -1,1\} \) at every point of the cube.
Let \(n\in \mathbb N\), \(I=\{ 0,\ldots ,n-1\} \), and \(f:\{ 0,1\} ^I\to \mathbb R\). Then its spectral one-norm is nonnegative:
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). For \(S\subseteq [n]\), let \(x^S\) be the bit vector obtained by flipping exactly the coordinates in \(S\). For all \(S,T\subseteq [n]\) and \(x\in \{ 0,1\} ^n\), \(\chi _T(x^S)=(-1)^{|T\cap S|}\chi _T(x)\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). For \(S\subseteq [n]\), let \(x^S\) be the bit vector obtained by flipping exactly the coordinates in \(S\). For every real cube function \(f\) and sets \(S,T\subseteq [n]\), \(\widehat{(x\mapsto f(x^S))}(T)=(-1)^{|T\cap S|}\widehat f(T)\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). For \(S\subseteq [n]\), let \(x^S\) be the bit vector obtained by flipping exactly the coordinates in \(S\). For every real cube function \(f\) and \(S\subseteq [n]\), \(\mathbb E[f(x)f(x^S)]=\sum _{T\subseteq [n]}(-1)^{|T\cap S|}\widehat f(T)^2\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), let \(x^S\) be the bit vector obtained by flipping exactly the coordinates in \(S\). For a real cube function \(f\) and \(S\subseteq [n]\), define \(d_S(f)=\tfrac 14\mathbb E[(f(x)-f(x^S))^2]\). For sign-valued \(f\), this equals \(\Pr [f(x)\ne f(x^S)]\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). For \(S\subseteq [n]\), let \(x^S\) be the bit vector obtained by flipping exactly the coordinates in \(S\). If \(f\colon \{ 0,1\} ^n\to \{ -1,1\} \), then \(\Pr [f(x)\ne f(x^S)]=(1-\sum _{T\subseteq [n]}(-1)^{|T\cap S|}\widehat f(T)^2)/2\) for each \(S\subseteq [n]\); equivalently, this is \(\tfrac 14\mathbb E[(f(x)-f(x^S))^2]\).
Let \(n,m\ge 0\), \(j_0\in [m]\), \(T\subseteq [n]\), and \(\pi \colon [n]\to [m]\), and put \(F_{j_0}(\pi )=\{ i:\pi (i)=j_0\} \). Then \((-1)^{|T\cap F_{j_0}(\pi )|}=\prod _{i=1}^n b_i\), where \(b_i=-1\) if \(i\in T\) and \(\pi (i)=j_0\), and \(b_i=1\) otherwise.
Let \(n\ge 0\), \(m\gt 0\) be integers, \(j_0\in [m]\), and \(T\subseteq [n]\). For \(F_{j_0}(\pi )=\{ i\in [n]:\pi (i)=j_0\} \), the sum over all colorings satisfies \(\sum _{\pi :[n]\to [m]}(-1)^{|T\cap F_{j_0}(\pi )|}=m^n(1-2/m)^{|T|}\).
Let \(n\in \mathbb N\) and let \(f:\{ 0,1\} ^n\to \mathbb R\) be a real-valued cube function, with uniform expectation \(\mathbb E[h]=2^{-n}\sum _{x\in \{ 0,1\} ^n}h(x)\) and Fourier coefficients \(\widehat f(S)=\mathbb E[f(x)\prod _{i\in S}(-1)^{x_i}]\) for \(S\subseteq \{ 0,\ldots ,n-1\} \). The degree-one Fourier weight \(W^1[f]=\sum _{|S|=1}\widehat f(S)^2\) satisfies
Let \(n\in \mathbb N\), let \(f:\{ 0,1\} ^n\to \mathbb R\) be a real-valued cube function, and let \((b_i)_{i\lt n}\) be real coefficients. With uniform expectation \(\mathbb E[h]=2^{-n}\sum _{x\in \{ 0,1\} ^n}h(x)\) and singleton Fourier coefficients \(\widehat f(\{ i\} )=\mathbb E[f(x)(-1)^{x_i}]\), one has
Let \(n\in \mathbb N\) and let \(g:\{ 0,1\} ^n\to \mathbb R\) be a real-valued cube function; write \(\mathbb E[h]=2^{-n}\sum _{x\in \{ 0,1\} ^n}h(x)\), \(\widehat g(\{ i\} )=\mathbb E[g(x)(-1)^{x_i}]\), and \(W^1[g]=\sum _{i\lt n}\widehat g(\{ i\} )^2\) for its degree-one Fourier weight. If \(W^1[g]\gt 0\), there exist real coefficients \((a_i)_{i\lt n}\) such that
Let \(n\in \mathbb N\), let \(J\subseteq \{ 0,\ldots ,n-1\} \) be nonempty, let \(b\in \{ 0,1\} ^n\), and put \(k=|J|\) and \(q(x)=\mathbf1_{\{ x_i=b_i\ \text{for every }i\in J\} }\) for \(x\in \{ 0,1\} ^n\). Using uniform expectation \(\mathbb E[h]=2^{-n}\sum _xh(x)\), singleton Fourier coefficients \(\widehat q(\{ i\} )=\mathbb E[q(x)(-1)^{x_i}]\), and degree-one Fourier weight \(W^1[q]=\sum _{i\lt n}\widehat q(\{ i\} )^2\), one has
For \(t\in \mathbb R\) and \(n\in \mathbb N\), define the \(\{ 0,1\} \)-valued function
where the quotient at \(n=0\) is taken to be \(0\). Write \(\mathbb E_n[h]=2^{-n}\sum _{x\in \{ 0,1\} ^n}h(x)\), \(\widehat h(\{ i\} )=\mathbb E_n[h(x)(-1)^{x_i}]\), and \(W^1[h]=\sum _{i\lt n}\widehat h(\{ i\} )^2\) for the degree-one Fourier weight; let \(\varphi (u)=(2\pi )^{-1/2}e^{-u^2/2}\) and \(\overline\Phi (t)=\int _t^\infty \varphi (u)\, du\) be the standard Gaussian density and upper tail. For every real \(t\), both limits hold as \(n\to \infty \) through the natural numbers:
Let \(n\in \mathbb N\) and let \((a_i)_{i\lt n}\) be real coefficients with \(\sum _{i\lt n}a_i^2=1\). For the linear form \(L(x)=\sum _{i\lt n}a_i(-1)^{x_i}\) on \(\{ 0,1\} ^n\), with uniform expectation \(\mathbb E[h]=2^{-n}\sum _{x\in \{ 0,1\} ^n}h(x)\), every \(t\in \mathbb R\) satisfies
Let \(n\in \mathbb N\), let \((a_i)_{i\lt n}\) be real coefficients with \(\sum _{i\lt n}a_i^2=1\), and let \(s\in \mathbb R\) satisfy \(s\ge 1\). For \(L(x)=\sum _{i\lt n}a_i(-1)^{x_i}\) on \(\{ 0,1\} ^n\) and uniform expectation \(\mathbb E[h]=2^{-n}\sum _{x\in \{ 0,1\} ^n}h(x)\), one has
For \(n\in \mathbb N\) and a real-valued function \(h\) on \(\{ 0,1\} ^n\), write \(\mathbb E[h]=2^{-n}\sum _{x\in \{ 0,1\} ^n}h(x)\), \(\widehat h(\{ i\} )=\mathbb E[h(x)(-1)^{x_i}]\), and \(W^1[h]=\sum _{i\lt n}\widehat h(\{ i\} )^2\) for its degree-one Fourier weight. There exists a real constant \(C\ge 0\) such that, for every \(n\in \mathbb N\), every \(f:\{ 0,1\} ^n\to \mathbb R\) satisfying \(f(x)\in \{ 0,1\} \) for every \(x\), and every real \(\alpha \) with \(\mathbb E[f]=\alpha \) and \(0\lt \alpha \le \tfrac 12\), one has
where \(\log \) is the natural logarithm.
For \(n\in \mathbb N\) and a real-valued function \(f\) on \(\{ 0,1\} ^n\), use uniform expectation \(\mathbb E[h]=2^{-n}\sum _xh(x)\), singleton Fourier coefficients \(\widehat f(\{ i\} )=\mathbb E[f(x)(-1)^{x_i}]\), degree-one Fourier weight \(W^1[f]=\sum _{i\lt n}\widehat f(\{ i\} )^2\), and degree-one part \(L_f(x)=\sum _{i\lt n}\widehat f(\{ i\} )(-1)^{x_i}\); define \(\tau (r)=1\) for \(r\ge 0\) and \(\tau (r)=-1\) for \(r\lt 0\). There exists a real constant \(C\ge 0\) such that, for every \(n\in \mathbb N\), every sign-valued function \(f:\{ 0,1\} ^n\to \{ -1,1\} \), and every real \(\varepsilon \ge 0\) with \(|\widehat f(\{ i\} )|\le \varepsilon \) for every \(i\lt n\), one has
where the probability is uniform on the cube.
For \(C\in \mathbb R\), define the linear FKN closeness property \(\mathcal F(C)\) as follows: for every \(n\in \mathbb N\) with \(n\gt 0\), every sign-valued function \(f:\{ 0,1\} ^n\to \{ -1,1\} \), and every real \(\delta \) with \(0\le \delta \le 1\) and \(1-\delta \le W^1[f]\), there exist an index \(i\lt n\) and a sign \(s\in \{ 1,-1\} \) such that
Here \(d_i(x)=(-1)^{x_i}\) is the \(i\)th dictator, the probability is uniform on \(\{ 0,1\} ^n\), and \(W^1[f]=\sum _{i\lt n}\widehat f(\{ i\} )^2\) is the degree-one Fourier weight with \(\widehat f(\{ i\} )=2^{-n}\sum _{x\in \{ 0,1\} ^n}f(x)(-1)^{x_i}\).
Let \(n\in \mathbb N\) and let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be sign-valued, with uniform expectation \(\mathbb E[h]=2^{-n}\sum _{x\in \{ 0,1\} ^n}h(x)\), singleton Fourier coefficients \(\widehat f(\{ i\} )=\mathbb E[f(x)(-1)^{x_i}]\), and degree-one Fourier weight \(W^1[f]=\sum _{i\lt n}\widehat f(\{ i\} )^2\). For every real \(\delta \) satisfying \(0\le 1-\delta \) and \(1-\delta \le |\mathbb E[f]|\), one has
The logarithm is natural, and at \(\delta =0\) both \(2/\delta \) and \(\log (2/\delta )\) are interpreted as \(0\).
For any dimension \(d\in \mathbb N\) and real-valued function \(h\) on \(\{ 0,1\} ^d\), set \(\widehat h(\{ j\} )=2^{-d}\sum _{x\in \{ 0,1\} ^d}h(x)(-1)^{x_j}\) and \(W^1[h]=\sum _{j\lt d}\widehat h(\{ j\} )^2\) for its singleton Fourier coefficients and degree-one Fourier weight, and let every probability on a cube be uniform. Let \(C\in \mathbb R\) satisfy \(C\ge 1\), and assume that for every \(m\in \mathbb N\) with \(m\gt 0\), every \(g:\{ 0,1\} ^m\to \{ -1,1\} \), and every real \(\eta \) with \(0\le \eta \le 1\) and \(1-\eta \le W^1[g]\), there exist \(j\lt m\) and \(u\in \{ 1,-1\} \) with \(\Pr [g(x)\ne u(-1)^{x_j}]\le C\eta \). Then for every \(n\in \mathbb N\) with \(n\gt 0\), every sign-valued \(f:\{ 0,1\} ^n\to \{ -1,1\} \), and every real \(\delta \) with \(0\le \delta \le 1\) and \(1-\delta \le W^1[f]\), there exist \(i\lt n\) and \(s\in \{ 1,-1\} \) such that
where \(\log \) is natural and, at \(\delta =0\), \(1/(C\delta )\) and its logarithm are interpreted as \(0\).
Let \(n,k\in \mathbb N\), let \(f,g:\{ 0,1\} ^n\to \{ -1,1\} \), and define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb E_x[h(x)\chi _S(x)]\) using uniform expectation, for \(S\subseteq \{ 1,\ldots ,n\} \) and \(h:\{ 0,1\} ^n\to \mathbb R\). Suppose \(f(x)=\operatorname {sgn}(p(x))\) for a real multilinear polynomial \(p(x)=\sum _{|S|\le k}a_S\chi _S(x)\), where \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). If \(\widehat g(S)=\widehat f(S)\) for every \(S\) with \(|S|\le k\), then \(g=f\).
Let \(n\in \mathbb N\) and \(f,g:\{ 0,1\} ^n\to \{ -1,1\} \), and define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb E_x[h(x)\chi _S(x)]\) for uniform \(x\), subsets \(S\subseteq \{ 1,\ldots ,n\} \), and \(h:\{ 0,1\} ^n\to \mathbb R\). Suppose \(f(x)=\operatorname {sgn}(a_\varnothing +\sum _{i=1}^na_i(-1)^{x_i})\) for real coefficients, where \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). If \(\widehat g(S)=\widehat f(S)\) whenever \(|S|\le 1\), then \(g=f\).
Let \(n\in \mathbb N\), \(f:\{ 0,1\} ^n\to \{ -1,1\} \), and \(0\lt \delta \le \tfrac 12\), and define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) for uniform \(x\) and \(S\subseteq \{ 1,\ldots ,n\} \). If \(y\) is obtained from \(x\) by flipping each bit independently with probability \(\delta \), write \(\operatorname {NS}_\delta (f)=\Pr [f(x)\ne f(y)]\); then
Let \(n\in \mathbb N\), \(f:\{ 0,1\} ^n\to \{ -1,1\} \), and \(0\lt \delta \le \tfrac 12\); write \(\operatorname {NS}_\delta (f)=\Pr [f(x)\ne f(y)]\), where \(x\) is uniform and \(y\) is obtained by independently flipping each bit of \(x\) with probability \(\delta \). There exists \(g:\{ 0,1\} ^n\to \{ -1,1\} \) of the form \(g(x)=\operatorname {sgn}(p(x))\), where \(p\) is a real multilinear polynomial of degree at most \(\lfloor 1/\delta \rfloor \), with \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\), such that
Let \(n,k\in \mathbb N\), define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) on \(\{ 0,1\} ^n\) for \(S\subseteq \{ 1,\ldots ,n\} \), and let \(p(x)=\sum _Sa_S\chi _S(x)\) have real coefficients with \(a_S=0\) whenever \(|S|\gt k\). Suppose \(f:\{ 0,1\} ^n\to \mathbb R\) satisfies \(f(x)=\operatorname {sgn}(p(x))\), where \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\), and put \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) using uniform expectation. Then
Let \(n,k\in \mathbb N\), and let \(f:\{ 0,1\} ^n\to \mathbb R\) satisfy \(f(x)=\operatorname {sgn}(0)=1\) for every \(x\), with the convention \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). For \(S\subseteq \{ 1,\ldots ,n\} \) define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) using uniform expectation; then
Let \(n,k\in \mathbb N\), \(c\in \mathbb R\) with \(c\le 1\), and \(f:\{ 0,1\} ^n\to \mathbb R\), and define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) for uniform \(x\) and \(S\subseteq \{ 1,\ldots ,n\} \). Suppose \(f\) is the sign of some real multilinear polynomial of degree at most \(k\), with \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\), and assume that every real polynomial \(p(x)=\sum _{|S|\le k}a_S\chi _S(x)\) satisfies
Then \(c\le \sum _{|S|\le k}\widehat f(S)^2\).
Let \(n\in \mathbb N\) and \(f:\{ 0,1\} ^n\to \{ -1,1\} \) satisfy \(f(x)=\operatorname {sgn}(a_\varnothing +\sum _{i=1}^na_i(-1)^{x_i})\) for real coefficients, where \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). Define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) for uniform \(x\) and \(S\subseteq \{ 1,\ldots ,n\} \); then
Let \(n,k\in \mathbb N\), \(f:\{ 0,1\} ^n\to \{ -1,1\} \), and \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) for \(S\subseteq \{ 1,\ldots ,n\} \). Suppose \(f(x)=\operatorname {sgn}(\sum _{|S|\le k}a_S\chi _S(x))\) for real coefficients, where \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\), and define \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) using uniform expectation. Then
Let \(n\in \mathbb N\), \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) on \(\{ 0,1\} ^n\), and \(p(x)=\sum _Sa_S\chi _S(x)\) be a nonzero real multilinear polynomial; let \(F\) be a finite family of subsets of \(\{ 1,\ldots ,n\} \) containing every \(S\) with \(a_S\ne 0\). Suppose \(f:\{ 0,1\} ^n\to \mathbb R\) satisfies \(f(x)=\operatorname {sgn}(p(x))\), where \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\), and define \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) using uniform expectation. Then \(1\le \sum _{S\in F}|\widehat f(S)|\).
Let \(n\in \mathbb N\), \(f:\{ 0,1\} ^n\to \{ -1,1\} \), \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\), and \(p(x)=\sum _Sa_S\chi _S(x)\) be a nonzero real multilinear polynomial on \(\{ 0,1\} ^n\). Suppose \(f(x)=\operatorname {sgn}(p(x))\), with \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\), and let \(F\) be a finite family of subsets of \(\{ 1,\ldots ,n\} \) satisfying \(a_S\ne 0\Rightarrow S\in F\). Writing \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) for uniform \(x\), one has \(1\le \sum _{S\in F}|\widehat f(S)|\).
Let \(n\in \mathbb N\) and define \(I:\{ 0,1\} ^{2n}\to \{ -1,1\} \) by \(I(u,v)=(-1)^{\sum _{i=1}^nu_iv_i}\) for \(u,v\in \{ 0,1\} ^n\). For \(S\subseteq \{ 1,\ldots ,2n\} \) define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat I(S)=\mathbb E_x[I(x)\chi _S(x)]\) using uniform expectation; every such \(S\) satisfies
Let \(n\in \mathbb N\) and define \(I(u,v)=(-1)^{\sum _{i=1}^nu_iv_i}\) for \(u,v\in \{ 0,1\} ^n\). Suppose the nonzero real multilinear polynomial \(p(x)=\sum _{S\subseteq \{ 1,\ldots ,2n\} }a_S\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\), satisfies \(I(x)=\operatorname {sgn}(p(x))\) for all \(x\in \{ 0,1\} ^{2n}\), with \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). Then \(p\) has at least \(2^n\) nonzero coefficients: \(2^n\le |\{ S:a_S\ne 0\} |\).
Let \(n\in \mathbb N\) with \(n\gt 0\), let \(f:\{ 0,1\} ^n\to \mathbb R\) be nonzero, and define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\), \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) for uniform \(x\), and \(L=\sum _{S\subseteq \{ 1,\ldots ,n\} }|\widehat f(S)|\). For every \(\delta \gt 0\) and \(s\in \mathbb N\) with \(4nL^2/\delta ^2\le s\), there exists a real multilinear polynomial \(q(x)=\sum _Sa_S\chi _S(x)\) satisfying
Let \(n\in \mathbb N\) with \(n\gt 0\), let \(f:\{ 0,1\} ^n\to \{ -1,1\} \), define \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat f(S)=\mathbb E_x[f(x)\chi _S(x)]\) for uniform \(x\) and \(S\subseteq \{ 1,\ldots ,n\} \), and set \(L=\sum _S|\widehat f(S)|\) and \(s=\lceil 4nL^2\rceil \). There exists a real multilinear polynomial \(p(x)=\sum _Sa_S\chi _S(x)\) with at most \(s\) nonzero coefficients such that \(f(x)=\operatorname {sgn}(p(x))\), where \(\operatorname {sgn}(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). Moreover, there exist subsets \(S_1,\ldots ,S_s\subseteq \{ 1,\ldots ,n\} \), with repetitions allowed, and signs \(\varepsilon _1,\ldots ,\varepsilon _s\in \{ -1,1\} \) such that
For measures \(\mu ,\nu \) on \(\mathbb R\), define \(D_K(\mu ,\nu )=\sup _{t\in \mathbb R}|\mu (({-\infty },t])-\nu (({-\infty },t])|\), taking the real value of each measure mass, assigning real value zero to infinite mass, and assigning supremum zero if the displayed set is unbounded. For probability measures this is the usual Kolmogorov distance.
There exists a universal real constant \(c\ge 0\) such that, for every \(n\ge 0\) and family of independent measurable real random variables \(X_1,\ldots ,X_n\) on any probability space, if each \(|X_i|^3\) is integrable, \(\mathbb E X_i=0\), and \(\sum _i\mathbb E X_i^2=1\), then \(D_K(\mathcal L(\sum _iX_i),\gamma )\le c\sum _i\mathbb E|X_i|^3\). Here \(\gamma \) is standard Gaussian measure and \(D_K(\mu ,\nu )=\sup _{t\in \mathbb R}|\mu (({-\infty },t])-\nu (({-\infty },t])|\) for probability measures.
There exists a universal real constant \(C\ge 0\) such that for every \(n\ge 0\), real coefficients \(a_1,\ldots ,a_n\), and real \(\varepsilon \), if \(\sum _i a_i^2=1\) and \(|a_i|\le \varepsilon \) for all \(i\), then \(|\mathbb E|\sum _i a_i\xi _i|-\sqrt{2/\pi }|\le C\varepsilon \), where the \(\xi _i\) are independent uniform signs.
Let \(G,H\) be independent standard real Gaussians. For every real \(\rho \), define \(Q(\rho )=\Pr [G\le 0,\ \rho G+\sqrt{\max (1-\rho ^2,0)}\, H\le 0]\); when \(|\rho |\le 1\), the two displayed coordinates form a standard Gaussian pair of correlation \(\rho \).
There exists a universal real constant \(C\ge 0\) with the following property. For every \(n\ge 0\), real coefficients \(a_i\), and real \(\varepsilon ,\rho \) with \(\sum _i a_i^2=1\), \(|a_i|\le \varepsilon \) for every \(i\), \(|\rho |\lt 1\), and \(\sum _i a_i(-1)^{x_i}\ne 0\) for every \(x\in \{ 0,1\} ^n\), let \(f(x)\) equal \(1\) when \(\sum _i a_i(-1)^{x_i}\ge 0\) and \(-1\) otherwise. Then \(|\operatorname {Stab}_\rho (f)-(2/\pi )\arcsin \rho |\le C\sqrt{\max (\varepsilon ,0)}/(1-\rho )\). For a real function \(h\) on a uniform bit cube, put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\), \(\widehat h(S)=\mathbb E[h\chi _S]\), and \(\operatorname {Stab}_\rho (h)=\sum _S\rho ^{|S|}\widehat h(S)^2\).
There exists a universal real constant \(C\ge 0\) such that for every \(0\le \rho \lt 1\), the majority function \(M_k\) on \(2k+1\) uniform bits, equal to \(1\) when \(\sum _i(-1)^{x_i}\gt 0\) and \(-1\) otherwise, satisfies \(\operatorname {Stab}_\rho (M_{k+1})\le \operatorname {Stab}_\rho (M_k)\) for all \(k\ge 0\), and \(\operatorname {Stab}_\rho (M_k)\to (2/\pi )\arcsin \rho \) as \(k\to \infty \). For every integer \(k\ge 0\), \((2/\pi )\arcsin \rho \le \operatorname {Stab}_\rho (M_k)\le (2/\pi )\arcsin \rho +C/(\sqrt{1-\rho ^2}\sqrt{2k+1})\). For a real function \(h\) on a uniform bit cube, put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\), \(\widehat h(S)=\mathbb E[h\chi _S]\), and \(\operatorname {Stab}_\rho (h)=\sum _S\rho ^{|S|}\widehat h(S)^2\).
For every \(0\lt \rho \lt 1\) and \(\varepsilon \gt 0\), there exists \(\tau \gt 0\) such that for all \(n\ge 0\) and \(f\colon \{ 0,1\} ^n\to [-1,1]\) with uniform mean zero and \(I_i(f)\le \tau \) for every \(i\), one has \(\operatorname {Stab}_\rho (f)\le (2/\pi )\arcsin \rho +\varepsilon \). Here \(I_i(f)=\mathbb E[(f(x)-f(x^{\{ i\} }))^2]/4\), where \(x^{\{ i\} }\) flips bit \(i\). For a real function \(h\) on a uniform bit cube, put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\), \(\widehat h(S)=\mathbb E[h\chi _S]\), and \(\operatorname {Stab}_\rho (h)=\sum _S\rho ^{|S|}\widehat h(S)^2\).
Let \(\gamma \) be the standard real Gaussian probability measure. For every measurable \(A\subseteq \mathbb R^2\), \((\gamma \otimes \gamma )(A)=\int _A(2\pi )^{-1}e^{-(x^2+y^2)/2}\, dx\, dy\), where the integral is nonnegative and extended-valued.
Let \(A\subseteq \mathbb R^2\) and \(\Theta \subseteq \mathbb R\) be measurable sets such that for all real \(r\gt 0\) and \(\theta \), \((r\cos \theta ,r\sin \theta )\in A\) if and only if \(\theta \in \Theta \). Then \(\int _A(2\pi )^{-1}e^{-(x^2+y^2)/2}\, dx\, dy=(2\pi )^{-1}\operatorname {Leb}(\Theta \cap (-\pi ,\pi ))\), with the integral and angular Lebesgue measure in the nonnegative extended reals.
The standard Gaussian probability measure \(\gamma \) on \(\mathbb R\) is the measure with Lebesgue density \(x\mapsto e^{-x^2/2}/\sqrt{2\pi }\). For every Borel set \(A\subseteq \mathbb R\), it satisfies
Let \(\gamma \) be the standard Gaussian probability measure on \(\mathbb R\), with density \(x\mapsto e^{-x^2/2}/\sqrt{2\pi }\). Its cumulative distribution function is the real-valued function defined for every \(t\in \mathbb R\) by \(\Phi (t)=\gamma (({-\infty },t])\).
Let \(\Phi (t)=\int _{-\infty }^{t} e^{-x^2/2}/\sqrt{2\pi }\, dx\) be the standard Gaussian cumulative distribution function. Its complementary tail function is defined for every \(t\in \mathbb R\) by \(\overline\Phi (t)=1-\Phi (t)\).
Let \(\Phi (t)=\int _{-\infty }^{t}e^{-x^2/2}/\sqrt{2\pi }\, dx\) be the standard Gaussian cumulative distribution function. Define \(Q:\mathbb R\to \mathbb R\) by
where the infimum is assigned the value \(0\) when its set is empty or unbounded below. Thus \(Q\) is the usual lower Gaussian quantile for \(0\lt \alpha \lt 1\), and \(Q(\alpha )=0\) when \(\alpha \le 0\) or \(\alpha \ge 1\).
Let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\) and \(\Phi (t)=\int _{-\infty }^{t}\varphi (x)\, dx\), and define \(Q(\alpha )=\inf \{ t\in \mathbb R:\alpha \le \Phi (t)\} \) for \(0\lt \alpha \lt 1\) and \(Q(\alpha )=0\) otherwise. The Gaussian isoperimetric profile is the function \(U:\mathbb R\to \mathbb R\) given by \(U(\alpha )=\varphi (Q(\alpha ))\). In particular its values at both endpoints and outside \([0,1]\) are \(U(\alpha )=1/\sqrt{2\pi }\) for \(\alpha \le 0\) or \(\alpha \ge 1\).
Let \(\gamma \) be the standard Gaussian probability measure on \(\mathbb R\), let \(\Phi (t)=\gamma (({-\infty },t])\), and let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\). For every \(t\in \mathbb R\), the cumulative probability equals the Lebesgue integral of the density:
Let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\) and \(\Phi (t)=\int _{-\infty }^{t}\varphi (x)\, dx\). For any real \(a,b\), with no ordering assumption, one has
where the interval integral is signed, so that \(\int _a^b=-\int _b^a\) when \(b\lt a\).
Let \(\Phi (t)=\int _{-\infty }^{t}e^{-x^2/2}/\sqrt{2\pi }\, dx\) be the standard Gaussian cumulative distribution function. It is strictly increasing: for every \(a,b\in \mathbb R\) with \(a\lt b\), one has \(\Phi (a)\lt \Phi (b)\).
Let \(\Phi (t)=\int _{-\infty }^{t}e^{-x^2/2}/\sqrt{2\pi }\, dx\) be the standard Gaussian cumulative distribution function. For every real \(\alpha \) satisfying \(0\lt \alpha \lt 1\), there exists a real \(t\) with \(\Phi (t)=\alpha \).
Let \(\Phi (t)=\int _{-\infty }^{t}e^{-x^2/2}/\sqrt{2\pi }\, dx\), and let \(Q(\alpha )\) be \(\inf \{ s\in \mathbb R:\alpha \le \Phi (s)\} \) for \(0\lt \alpha \lt 1\) and \(0\) otherwise. For real \(\alpha ,t\) satisfying \(\Phi (t)=\alpha \), one has \(Q(\alpha )=t\).
Let \(\Phi (t)=\int _{-\infty }^{t}e^{-x^2/2}/\sqrt{2\pi }\, dx\), and let \(Q(\alpha )\) be \(\inf \{ t\in \mathbb R:\alpha \le \Phi (t)\} \) for \(0\lt \alpha \lt 1\) and \(0\) otherwise. For every \(\alpha \in \mathbb R\) with \(0\lt \alpha \lt 1\), one has \(\Phi (Q(\alpha ))=\alpha \).
Let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\), let \(F(0)=0\) and \(F(x)=\varphi (x)(x^{-3}-x^{-1})\) for \(x\ne 0\), and let \(H(0)=\varphi (0)\) and \(H(x)=\varphi (x)(1-3x^{-4})\) for \(x\ne 0\). For every nonzero real \(x\), the function \(F\) is differentiable at \(x\) with derivative \(F'(x)=H(x)\).
Let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\), and define \(H:\mathbb R\to \mathbb R\) by \(H(0)=\varphi (0)\) and \(H(x)=\varphi (x)(1-3x^{-4})\) for \(x\ne 0\). For every real \(t\lt 0\), the function \(H\) is integrable on \((-\infty ,t]\) with respect to Lebesgue measure.
Let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\) and \(\Phi (t)=\int _{-\infty }^{t}\varphi (x)\, dx\). Define the real-valued function \(R\) for every \(t\in \mathbb R\) by
On \(t\lt 0\), this is the lower-tail Mills ratio \(\Phi (t)|t|/\varphi (t)\).
Let \(\Phi (t)=\int _{-\infty }^{t}e^{-x^2/2}/\sqrt{2\pi }\, dx\) be the standard Gaussian cumulative distribution function. Define \(L:\mathbb R\to \mathbb R\) for every real \(t\) by
where \(\log \) is the natural logarithm.
Let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\), \(\Phi (t)=\int _{-\infty }^{t}\varphi (x)\, dx\), \(R(t)=\Phi (t)(-t)/\varphi (t)\), and \(L(t)=2\log (1/\Phi (t))\). For every real \(t\lt 0\), with natural logarithms, one has
Let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\), \(\Phi (t)=\int _{-\infty }^{t}\varphi (x)\, dx\), and \(L(t)=2\log (1/\Phi (t))\), with natural logarithm. Then
Let \(\Phi (t)=\int _{-\infty }^{t}e^{-x^2/2}/\sqrt{2\pi }\, dx\), and let \(Q(\alpha )=\inf \{ t\in \mathbb R:\alpha \le \Phi (t)\} \) for \(0\lt \alpha \lt 1\) and \(Q(\alpha )=0\) otherwise. As a real \(\alpha \) tends to zero through positive values, the quantile tends to negative infinity:
Let \(\varphi (x)=e^{-x^2/2}/\sqrt{2\pi }\), \(\Phi (t)=\int _{-\infty }^{t}\varphi (x)\, dx\), and \(Q(\alpha )=\inf \{ t\in \mathbb R:\alpha \le \Phi (t)\} \) for \(0\lt \alpha \lt 1\), with \(Q(\alpha )=0\) otherwise. Define \(U(\alpha )=\varphi (Q(\alpha ))\) for every real \(\alpha \), so \(U(\alpha )=1/\sqrt{2\pi }\) for \(\alpha \le 0\) or \(\alpha \ge 1\). As \(\alpha \to 0\) through positive values, this profile is asymptotically equivalent to \(\alpha \sqrt{2\log (1/\alpha )}\); equivalently, with natural logarithm,
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. For a real cube function \(h\), write \(D_i h(x)=\tfrac 12(h(x_{i\leftarrow 0})-h(x_{i\leftarrow 1}))\), where \(x_{i\leftarrow b}\) replaces bit \(i\) by \(b\). For every coordinate \(i\) and point \(x\), \(D_i h(x)=\tfrac 12(h(x)-h(x^{\{ i\} }))(-1)^{x_i}\), where \(x^{\{ i\} }\) flips bit \(i\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. For a real cube function \(h\), write \(D_i h(x)=\tfrac 12(h(x_{i\leftarrow 0})-h(x_{i\leftarrow 1}))\), where \(x_{i\leftarrow b}\) replaces bit \(i\) by \(b\). Set \(I_i(h)=\tfrac 14\mathbb E[(h(x)-h(x^{\{ i\} }))^2]\) and \(I(h)=\sum _i I_i(h)\), where \(x^{\{ i\} }\) flips bit \(i\). For every real cube function \(h\) and \(i\in [n]\), \(I_i(h)=\mathbb E[(D_i h)^2]\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Let \(g\) be a real cube function and \(i\in [n]\), and define \(E_i g(x)=\tfrac 12(g(x_{i\leftarrow 0})+g(x_{i\leftarrow 1}))\), where the indicated update replaces bit \(i\). Then \(\mathbb E[E_i g]=\mathbb E g\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). For a real cube function \(h\), write \(D_i h(x)=\tfrac 12(h(x_{i\leftarrow 0})-h(x_{i\leftarrow 1}))\), where \(x_{i\leftarrow b}\) replaces bit \(i\) by \(b\). For every real cube function \(h\) and coordinate \(i\), \(\widehat h(\{ i\} )=\mathbb E[D_i h]\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. For a real cube function \(h\), write \(D_i h(x)=\tfrac 12(h(x_{i\leftarrow 0})-h(x_{i\leftarrow 1}))\), where \(x_{i\leftarrow b}\) replaces bit \(i\) by \(b\). For any \(h\colon \{ 0,1\} ^n\to \{ -1,1\} \), coordinate \(i\), and cube point \(x\), \((D_i h(x))^2=|D_i h(x)|\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). For a real cube function \(h\), write \(D_i h(x)=\tfrac 12(h(x_{i\leftarrow 0})-h(x_{i\leftarrow 1}))\), where \(x_{i\leftarrow b}\) replaces bit \(i\) by \(b\). Set \(I_i(h)=\tfrac 14\mathbb E[(h(x)-h(x^{\{ i\} }))^2]\) and \(I(h)=\sum _i I_i(h)\), where \(x^{\{ i\} }\) flips bit \(i\). If \(h\) is sign-valued and either \(D_i h(x)\ge 0\) for all \(x\) or \(D_i h(x)\le 0\) for all \(x\), then \(I_i(h)=|\widehat h(\{ i\} )|\).
Let \(n\ge 0\). Let \(c\in \mathbb R\), \(a\in \mathbb R^n\), and let \(A_{c,a}(S)\) equal \(c\) for \(S=\varnothing \), \(a_i\) for \(S=\{ i\} \), and zero for all larger \(S\). This coefficient family has degree at most one: \(A_{c,a}(S)=0\) whenever \(|S|\gt 1\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). Write \(\operatorname {sgn}_+(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). If \(f\) is the sign \(\operatorname {sgn}_+(P)\) of a Walsh polynomial \(P\) whose coefficients vanish on sets larger than one, then there exist real \(c\) and \(a_1,\ldots ,a_n\) such that \(f(x)=\operatorname {sgn}_+(c+\sum _i a_i(-1)^{x_i})\) for every \(x\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Write \(\operatorname {sgn}_+(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). For every real \(c\) and \(a\in \mathbb R^n\), the function \(x\mapsto \operatorname {sgn}_+(c+\sum _i a_i(-1)^{x_i})\) is a linear threshold function, meaning the sign of a Walsh polynomial with no coefficient above degree one.
Let \(n\ge 0\), \(c\in \mathbb R\), \(a\in \mathbb R^n\), \(i\in [n]\), \(x\in \{ 0,1\} ^n\), and \(b\in \{ 0,1\} \). If \(x_{i\leftarrow b}\) replaces bit \(i\) by \(b\), then \(c+\sum _j a_j(-1)^{(x_{i\leftarrow b})_j}=(c+\sum _{j\ne i}a_j(-1)^{x_j})+a_i(-1)^b\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Write \(\operatorname {sgn}_+(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). For a real cube function \(h\), write \(D_i h(x)=\tfrac 12(h(x_{i\leftarrow 0})-h(x_{i\leftarrow 1}))\), where \(x_{i\leftarrow b}\) replaces bit \(i\) by \(b\). For any real \(c\), \(a\in \mathbb R^n\), and \(i\in [n]\), put \(f(x)=\operatorname {sgn}_+(c+\sum _j a_j(-1)^{x_j})\). Either \(D_i f(x)\ge 0\) for every \(x\), or \(D_i f(x)\le 0\) for every \(x\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). Every \(f\colon \{ 0,1\} ^n\to \{ -1,1\} \) satisfies \(\sum _{i=1}^n\widehat f(\{ i\} )^2\le 1\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Write \(\operatorname {sgn}_+(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). Set \(I_i(h)=\tfrac 14\mathbb E[(h(x)-h(x^{\{ i\} }))^2]\) and \(I(h)=\sum _i I_i(h)\), where \(x^{\{ i\} }\) flips bit \(i\). If \(f(x)=\operatorname {sgn}_+(c+\sum _i a_i(-1)^{x_i})\) for some real \(c,a_i\), then \(I(f)\le \sqrt n\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients; degree at most \(k\) means \(a_S=0\) for \(|S|\gt k\). If \(P\) has degree at most the integer \(k\ge 0\) and \(0\le \rho \le 1\), define \(T_\rho P=\sum _S\rho ^{|S|}a_S\chi _S\). Then \(\rho ^{2k}\sum _Sa_S^2\le \mathbb E[(T_\rho P)^2]\), with natural powers including \(0^0=1\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients; degree at most \(k\) means \(a_S=0\) for \(|S|\gt k\). If \(P\) has degree at most the integer \(k\ge 0\) and \(1\lt t\lt 2\), then \((t-1)^k\sum _Sa_S^2\le (\mathbb E|P|^t)^{2/t}\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients; degree at most \(k\) means \(a_S=0\) for \(|S|\gt k\). If \(P\) has degree at most the integer \(k\ge 0\), then \(e^{-2k}\sum _{S\subseteq [n]}a_S^2\le (\mathbb E|P|)^2\).
Let \(m\in \mathbb N\), put \(n=2m+1\), and define \(M_m:\{ \pm 1\} ^n\to \mathbb R\) by \(M_m(x)=1\) when more than \(m\) coordinates equal \(1\), and \(M_m(x)=-1\) otherwise. Then \(M_m(-x)=-M_m(x)\) for every \(x\in \{ \pm 1\} ^n\), where \(-x\) negates every coordinate.
Let \(n\in \mathbb N\) and \(f:\{ \pm 1\} ^n\to \mathbb R\); for \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat f(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}f(x)\prod _{i\in S}x_i\), and for every \(k\in \mathbb N\) define \(W_k(f)=\sum _{|S|=k}\widehat f(S)^2\). Then \(W_k(f)\ge 0\) for every \(k\in \mathbb N\).
Let \(n\in \mathbb N\) and \(f:\{ \pm 1\} ^n\to \mathbb R\); for \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat f(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}f(x)\prod _{i\in S}x_i\), and for every \(k\in \mathbb N\) define \(W_k(f)=\sum _{|S|=k}\widehat f(S)^2\). Then \(W_k(f)=0\) for every \(k\in \mathbb N\) with \(n\lt k\).
Let \(n\in \mathbb N\) and \(f:\{ \pm 1\} ^n\to \mathbb R\); for \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat f(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}f(x)\prod _{i\in S}x_i\), and for every \(k\in \mathbb N\) define \(W_k(f)=\sum _{|S|=k}\widehat f(S)^2\). For every \(g:\mathbb N\to \mathbb R\),
Let \(n\in \mathbb N\), \(\rho \in \mathbb R\), and \(f:\{ \pm 1\} ^n\to \mathbb R\); for \(S\subseteq \{ 1,\ldots ,n\} \) define \(\chi _S(x)=\prod _{i\in S}x_i\), \(\widehat f(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}f(x)\chi _S(x)\), \(W_k(f)=\sum _{|S|=k}\widehat f(S)^2\) for \(k\in \mathbb N\), and \(T_\rho f(x)=\sum _{S\subseteq \{ 1,\ldots ,n\} }\rho ^{|S|}\widehat f(S)\chi _S(x)\). The noise stability \(\operatorname {Stab}_\rho (f)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}f(x)T_\rho f(x)\) satisfies
Let \(n\in \mathbb N\) and \(f:\{ \pm 1\} ^n\to \mathbb R\) satisfy \(f(x)\in \{ 1,-1\} \) for every \(x\); for \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat f(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}f(x)\prod _{i\in S}x_i\) and \(W_k(f)=\sum _{|S|=k}\widehat f(S)^2\) for \(k\in \mathbb N\). Then \(\sum _{k=0}^n W_k(f)=1\).
Let \(N\in \mathbb N\), \(\rho \in [0,1]\), and \(a,b:\mathbb N\to \mathbb R\) satisfy \(\sum _{k=0}^{j-1}b(k)\le \sum _{k=0}^{j-1}a(k)\) for every integer \(0\le j\le N\) and \(\sum _{k=0}^{N-1}b(k)=\sum _{k=0}^{N-1}a(k)\), with sums over no indices interpreted as zero. Then
Let \(m\in \mathbb N\), put \(n=2m+1\), and define \(M_m:\{ \pm 1\} ^n\to \mathbb R\) by \(M_m(x)=1\) when more than \(m\) coordinates equal \(1\), and \(M_m(x)=-1\) otherwise. For \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat M_m(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}M_m(x)\prod _{i\in S}x_i\). If \(S\subseteq \{ 1,\ldots ,n\} \) has even cardinality, then \(\widehat M_m(S)=0\).
Let \(m\in \mathbb N\), put \(n=2m+1\), and define \(M_m:\{ \pm 1\} ^n\to \mathbb R\) by \(M_m(x)=1\) when more than \(m\) coordinates equal \(1\), and \(M_m(x)=-1\) otherwise. For \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat M_m(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}M_m(x)\prod _{i\in S}x_i\). If \(S\subseteq \{ 1,\ldots ,n\} \) has cardinality \(2r+1\) for an integer \(r\ge 0\), then
Let \(m\in \mathbb N\), put \(n=2m+1\), and define \(M_m:\{ \pm 1\} ^n\to \mathbb R\) by \(M_m(x)=1\) when more than \(m\) coordinates equal \(1\), and \(M_m(x)=-1\) otherwise. For \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat M_m(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}M_m(x)\prod _{i\in S}x_i\). If \(S,T\subseteq \{ 1,\ldots ,n\} \) satisfy \(|S|+|T|=2m+2\), then
Let \(m\in \mathbb N\), put \(n=2m+1\), and define \(M_m:\{ \pm 1\} ^n\to \mathbb R\) by \(M_m(x)=1\) when more than \(m\) coordinates equal \(1\), and \(M_m(x)=-1\) otherwise. For \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat M_m(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}M_m(x)\prod _{i\in S}x_i\), and for every \(k\in \mathbb N\) define \(W_k(M_m)=\sum _{|S|=k}\widehat M_m(S)^2\). If \(k\in \mathbb N\) is even, then \(W_k(M_m)=0\).
Let \(m\in \mathbb N\), put \(n=2m+1\), and define \(M_m:\{ \pm 1\} ^n\to \mathbb R\) by \(M_m(x)=1\) when more than \(m\) coordinates equal \(1\), and \(M_m(x)=-1\) otherwise. For \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat M_m(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}M_m(x)\prod _{i\in S}x_i\), and for every \(k\in \mathbb N\) define \(W_k(M_m)=\sum _{|S|=k}\widehat M_m(S)^2\). For every odd \(k\in \mathbb N\), put \(r=(k-1)/2\); then
where \(\binom ab=0\) for \(b\gt a\) and a real quotient with zero denominator is interpreted as zero.
Let \(m\in \mathbb N\), put \(n=2m+1\), and define \(M_m:\{ \pm 1\} ^n\to \mathbb R\) by \(M_m(x)=1\) when more than \(m\) coordinates equal \(1\), and \(M_m(x)=-1\) otherwise. For \(S\subseteq \{ 1,\ldots ,n\} \) define \(\widehat M_m(S)=2^{-n}\sum _{x\in \{ \pm 1\} ^n}M_m(x)\prod _{i\in S}x_i\), and for every \(k\in \mathbb N\) define \(W_k(M_m)=\sum _{|S|=k}\widehat M_m(S)^2\). For every \(k\in \mathbb N\) with \(k\le 2m+1\),
For each \(q\in \mathbb N\), define \(M_q:\{ \pm 1\} ^{2q+1}\to \mathbb R\) to equal \(1\) when more than \(q\) coordinates equal \(1\), and \(-1\) otherwise; for \(S\subseteq \{ 1,\ldots ,2q+1\} \) define \(\widehat M_q(S)=2^{-(2q+1)}\sum _{x\in \{ \pm 1\} ^{2q+1}}M_q(x)\prod _{i\in S}x_i\), and for \(k\in \mathbb N\) define \(W_k(M_q)=\sum _{|S|=k}\widehat M_q(S)^2\). For every \(m,k\in \mathbb N\) with \(k\) odd and \(k\le 2m+1\), one has \(W_k(M_{m+1})\lt W_k(M_m)\).
For each \(q\in \mathbb N\), define \(M_q:\{ \pm 1\} ^{2q+1}\to \mathbb R\) to equal \(1\) when more than \(q\) coordinates equal \(1\), and \(-1\) otherwise; for \(S\subseteq \{ 1,\ldots ,2q+1\} \) define \(\widehat M_q(S)=2^{-(2q+1)}\sum _{x\in \{ \pm 1\} ^{2q+1}}M_q(x)\prod _{i\in S}x_i\), and for \(k\in \mathbb N\) define \(W_k(M_q)=\sum _{|S|=k}\widehat M_q(S)^2\). For each fixed odd integer \(k\ge 1\), set \(L_k=\frac{4}{\pi k2^k}\binom {k-1}{(k-1)/2}\); then \(\lim _{m\to \infty }W_k(M_m)=L_k\), and for every \(m\in \mathbb N\) satisfying \(k\le \lfloor (2m+1)/2\rfloor \) one has
For each \(q\in \mathbb N\), define \(M_q:\{ \pm 1\} ^{2q+1}\to \mathbb R\) to equal \(1\) when more than \(q\) coordinates equal \(1\), and \(-1\) otherwise; for \(S\subseteq \{ 1,\ldots ,2q+1\} \) define \(\widehat M_q(S)=2^{-(2q+1)}\sum _{x\in \{ \pm 1\} ^{2q+1}}M_q(x)\prod _{i\in S}x_i\), and for \(k\in \mathbb N\) define \(W_k(M_q)=\sum _{|S|=k}\widehat M_q(S)^2\). For every \(m,j\in \mathbb N\) satisfying \(j\le 2(m+1)+2\),
where both sums are zero when \(j=0\).
Let \(m\in \mathbb N\) and \(\rho \in [0,1]\), and for each \(q\in \mathbb N\) define \(M_q:\{ \pm 1\} ^{2q+1}\to \mathbb R\) to equal \(1\) when more than \(q\) coordinates equal \(1\), and \(-1\) otherwise. For any \(d\in \mathbb N\) and \(f:\{ \pm 1\} ^d\to \mathbb R\), define \(\operatorname {Stab}_\rho (f)=\mathbb E[f(X)f(Y)]\), where \(X\) is uniform on \(\{ \pm 1\} ^d\) and, independently in each coordinate, \(Y_i=-X_i\) with probability \((1-\rho )/2\) and \(Y_i=X_i\) otherwise; then
For each \(m\in \mathbb N\), define \(M_m:\{ \pm 1\} ^{2m+1}\to \mathbb R\) to equal \(1\) when more than \(m\) coordinates equal \(1\), and \(-1\) otherwise; for \(S\subseteq \{ 1,\ldots ,2m+1\} \) define \(\widehat M_m(S)=2^{-(2m+1)}\sum _{x\in \{ \pm 1\} ^{2m+1}}M_m(x)\prod _{i\in S}x_i\), \(W_k(M_m)=\sum _{|S|=k}\widehat M_m(S)^2\), and \(W_{\gt k}(M_m)=\sum _{|S|\gt k}\widehat M_m(S)^2\) for \(k\in \mathbb N\), and put \(A=(2/\pi )^{3/2}\). There exists one real constant \(C\ge 0\) such that for every \(m,k\in \mathbb N\) with \(k\) odd, \(k\gt 0\), and \(2k^2\le 2m+1\), both estimates hold:
For \(k,n\in \mathbb N\), the degree-at-most-\(k\) polynomial threshold class consists of functions \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) for which there is a real multilinear polynomial \(p\) of degree at most \(k\) such that \(f(x)=\operatorname {sgn}_{+}(p((-1)^{x_1},\ldots ,(-1)^{x_n}))\) for every \(x\), where \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(\operatorname {sgn}_{+}(t)=-1\) for \(t\lt 0\).
For every \(n\in \mathbb N\), the linear threshold class consists of functions \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) admitting real coefficients \(c,a_1,\ldots ,a_n\) with
where \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\).
For a family \(B=(B_n)_{n\in \mathbb N}\) of sets of real-valued functions on \(\{ 0,1\} ^{n}\), uniform noise stability means that every member is \(\{ -1,1\} \)-valued and there exists a function \(\varepsilon :\mathbb R\to \mathbb R\) with \(0\leq \varepsilon (\delta )\leq 1\) for \(0\leq \delta \leq 1/2\) and \(\lim _{\delta \to 0^{+}}\varepsilon (\delta )=0\), such that
Here \(X\) is uniform on \(\{ 0,1\} ^{n}\) and \(Y\) is obtained by flipping each coordinate of \(X\) independently with probability \(\delta \); no range condition is imposed on \(\varepsilon \) outside \([0,1/2]\).
For every \(n,m\in \mathbb N\), a function \(f:\{ 0,1\} ^{n}\to \mathbb R\), a vector \(z\in \{ 0,1\} ^{n}\) and a map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \), negating according to \(z\) and identifying according to \(\pi \) gives the function \(g:\{ 0,1\} ^{m}\to \mathbb R\) defined by
For a family \(B=(B_n)_{n\in \mathbb N}\) of sets of real-valued functions on \(\{ 0,1\} ^{n}\), closure under input negation and identification means that, for all \(n,m\in \mathbb N\), \(f\in B_n\), \(z\in \{ 0,1\} ^{n}\) and maps \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \), the function
belongs to \(B_m\).
For a real parameter \(\delta \), the block count is the natural number
where the floor is the integer floor and \(0^{-1}\) is defined to be \(0\); in particular \(m(\delta )=0\) for \(\delta \leq 0\).
For every \(n\in \mathbb N\), \(f:\{ 0,1\} ^{n}\to \mathbb R\) and real \(\rho \geq 0\), the noise stability
is nonnegative, where \(\widehat f(S)=2^{-n}\sum _{x\in \{ 0,1\} ^{n}}f(x)\prod _{i\in S}(-1)^{x_i}\) is the normalized Walsh coefficient (and \(\rho ^0=1\), including at \(\rho =0\)).
For every \(n\in \mathbb N\), real-valued \(f:\{ 0,1\} ^{n}\to \mathbb R\) and real \(\delta \leq 1/2\), including negative \(\delta \), the total real-valued noise sensitivity satisfies
where \(\widehat f(S)=2^{-n}\sum _{x\in \{ 0,1\} ^{n}}f(x)\prod _{i\in S}(-1)^{x_i}\).
Let \(B=(B_n)_{n\in \mathbb N}\) be a family of sets of functions \(\{ 0,1\} ^{n}\to \{ -1,1\} \) and let \(\varepsilon :\mathbb R\to \mathbb R\) satisfy \(\varepsilon (\delta )\geq 0\) for \(0\leq \delta \leq 1/2\), \(\lim _{\delta \to 0^{+}}\varepsilon (\delta )=0\) and \(\operatorname {NS}_{\delta }(f)\leq \varepsilon (\delta )\) for every \(n\), \(f\in B_n\) and \(0\lt \delta \leq 1/2\). Then \(B\) is uniformly noise-stable: there exists \(\eta :\mathbb R\to \mathbb R\) with \(0\leq \eta (\delta )\leq 1\) on \([0,1/2]\), \(\lim _{\delta \to 0^{+}}\eta (\delta )=0\) and \(\operatorname {NS}_{\delta }(f)\leq \eta (\delta )\) for every such \(n,f,\delta \), where \(\operatorname {NS}_{\delta }(f)=\Pr [f(X)\neq f(Y)]\), \(X\) is uniform on the cube and \(Y\) flips each coordinate independently with probability \(\delta \).
For every \(n\in \mathbb N\), every member \(f\) of the linear threshold class on \(\{ 0,1\} ^{n}\) takes values in \(\{ -1,1\} \); the class consists of functions \(f(x)=\operatorname {sgn}_{+}(c+\sum _{i=1}^{n}a_i(-1)^{x_i})\) with real coefficients, where \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\).
For all \(n,m\in \mathbb N\), real coefficients \(c,a_1,\ldots ,a_n\), \(z\in \{ 0,1\} ^{n}\) and maps \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \), if \(f(x)=\operatorname {sgn}_{+}(c+\sum _{i=1}^{n}a_i(-1)^{x_i})\), then \(g(w)=f((z_i+w_{\pi (i)}\pmod2)_{i=1}^{n})\) is a \(\{ -1,1\} \)-valued linear threshold function on \(\{ 0,1\} ^{m}\), meaning that \(g(w)=\operatorname {sgn}_{+}(b+\sum _{j=1}^{m}u_j(-1)^{w_j})\) for some real \(b,u_1,\ldots ,u_m\). Here \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\).
For \(n,m\in \mathbb N\), a map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \) and a subset \(S\subseteq \{ 1,\ldots ,n\} \), the odd-multiplicity image is the set
For all \(n,m\in \mathbb N\), maps \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \) and subsets \(S\subseteq \{ 1,\ldots ,n\} \), the set \(O_{\pi }(S)=\{ j\in \{ 1,\ldots ,m\} :|\{ i\in S:\pi (i)=j\} |\text{ is odd}\} \) satisfies \(|O_{\pi }(S)|\leq |S|\).
For \(n,m\in \mathbb N\), a map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \), \(S\subseteq \{ 1,\ldots ,n\} \) and \(w\in \{ 0,1\} ^{m}\), set \(O_{\pi }(S)=\{ j\in \{ 1,\ldots ,m\} :|\{ i\in S:\pi (i)=j\} |\text{ is odd}\} \); then
For \(n,m\in \mathbb N\), \(z\in \{ 0,1\} ^{n}\), \(w\in \{ 0,1\} ^{m}\), a map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \) and \(S\subseteq \{ 1,\ldots ,n\} \), define \(\chi _A(x)=\prod _{i\in A}(-1)^{x_i}\) for subsets of the coordinate set of \(x\), and set \(O_{\pi }(S)=\{ j\in \{ 1,\ldots ,m\} :|\{ i\in S:\pi (i)=j\} |\text{ is odd}\} \); then
For every \(k,n,m\in \mathbb N\), a function \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) of the form \(f(x)=\operatorname {sgn}_{+}(p((-1)^{x_1},\ldots ,(-1)^{x_n}))\) for a real multilinear polynomial \(p\) of degree at most \(k\), every \(z\in \{ 0,1\} ^{n}\) and every map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \), the function \(g(w)=f((z_i+w_{\pi (i)}\pmod2)_{i=1}^{n})\) is also \(\{ -1,1\} \)-valued and has a representation \(g(w)=\operatorname {sgn}_{+}(q((-1)^{w_1},\ldots ,(-1)^{w_m}))\) with a real multilinear polynomial \(q\) of degree at most \(k\). Here \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\).
For every \(n,k,k'\in \mathbb N\) with \(k\leq k'\), every \(\{ -1,1\} \)-valued function on \(\{ 0,1\} ^{n}\) represented by \(\operatorname {sgn}_{+}(p((-1)^{x_1},\ldots ,(-1)^{x_n}))\) for a real multilinear polynomial \(p\) of degree at most \(k\) also admits such a representation of degree at most \(k'\), where \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\).
For every \(n,m\in \mathbb N\), \(f:\{ 0,1\} ^{n}\to \mathbb R\), \(z\in \{ 0,1\} ^{n}\), \(w\in \{ 0,1\} ^{m}\) and map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \), define \(g(v)=f((u_i(v))_{i=1}^{n})\), where \(u_i(v)=1-v_{\pi (i)}\) when \(z_i=1\) and \(u_i(v)=v_{\pi (i)}\) when \(z_i=0\); then
where \(\oplus \) is coordinatewise addition modulo \(2\).
For \(n,m\in \mathbb N\), \(f:\{ 0,1\} ^{n}\to \mathbb R\), \(z\in \{ 0,1\} ^{n}\), a map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \), \(w\in \{ 0,1\} ^{m}\) and \(j\in \{ 1,\ldots ,m\} \), define \(g(v)=f(z\oplus (v_{\pi (i)})_{i=1}^{n})\), where \(\oplus \) is coordinatewise addition modulo \(2\), and write \(x^{A}\) for the vector obtained by flipping precisely the coordinates in \(A\); then
For \(n,m\in \mathbb N\), \(f:\{ 0,1\} ^{n}\to \mathbb R\), a map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \) and \(j\in \{ 1,\ldots ,m\} \), define \(g_z(w)=f(z\oplus (w_{\pi (i)})_{i=1}^{n})\), where \(\oplus \) is coordinatewise addition modulo \(2\), and write \(x^{A}\) for the vector obtained by flipping the coordinates in \(A\); then
Each expectation is uniform on the full Boolean cube of its variable’s dimension.
For every \(n\in \mathbb N\), positive integer \(m\), \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) and \(j\in \{ 1,\ldots ,m\} \),
Here \(X\) is uniform on \(\{ 0,1\} ^{n}\), \(X^{A}\) flips precisely the coordinates of \(A\), and \(\operatorname {NS}_{1/m}(f)\) is \(\Pr [f(X)\neq f(Y)]\) when each coordinate of \(X\) is independently flipped with probability \(1/m\) to obtain \(Y\).
For every \(n\in \mathbb N\), \(f:\{ 0,1\} ^{n}\to \mathbb R\) and real \(\delta \leq \delta '\leq 1/2\), with no lower bound on either rate, one has \(\operatorname {NS}_{\delta }(f)\leq \operatorname {NS}_{\delta '}(f)\), where the total real-valued extension is
Let \(B=(B_n)_{n\in \mathbb N}\) be a family of sets of functions \(\{ 0,1\} ^{n}\to \{ -1,1\} \), closed under all substitutions \(w\mapsto (z_i+w_{\pi (i)}\pmod2)_{i=1}^{n}\) for every \(n,m\in \mathbb N\), \(z\in \{ 0,1\} ^{n}\) and map \(\pi :\{ 1,\ldots ,n\} \to \{ 1,\ldots ,m\} \), and let \(A:\mathbb N\to \mathbb R\) satisfy
where \(X\) is uniform on \(\{ 0,1\} ^{n}\) and \(X^{\{ i\} }\) flips coordinate \(i\). Then for every \(0\lt \delta \leq 1/2\), \(n\in \mathbb N\) and \(f\in B_n\), setting \(m=\lfloor 1/\delta \rfloor \) gives
where \(\operatorname {NS}_{\delta }(f)=\Pr [f(X)\neq f(Y)]\) and \(Y\) independently flips each coordinate of \(X\) with probability \(\delta \).
There exists a universal real constant \(C\geq 0\) such that, for every \(n\in \mathbb N\), every linear threshold function \(f(x)=\operatorname {sgn}_{+}(c+\sum _{i=1}^{n}a_i(-1)^{x_i})\) on \(\{ 0,1\} ^{n}\) with real coefficients and every \(0\lt \delta \leq 1/2\),
Here \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\), and \(\operatorname {NS}_{\delta }(f)=\Pr [f(X)\neq f(Y)]\) with uniform \(X\) and independent coordinate flips of probability \(\delta \) producing \(Y\).
The family of linear threshold functions on Boolean cubes is uniformly noise-stable: every function has the form \(f(x)=\operatorname {sgn}_{+}(c+\sum _{i=1}^{n}a_i(-1)^{x_i})\) for real coefficients, is \(\{ -1,1\} \)-valued, and there exists \(\varepsilon :\mathbb R\to \mathbb R\) satisfying \(0\leq \varepsilon (\delta )\leq 1\) on \([0,1/2]\), \(\lim _{\delta \to 0^{+}}\varepsilon (\delta )=0\) and
for every dimension \(n\in \mathbb N\), every such \(f\) and every \(0\lt \delta \leq 1/2\). Here \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\), \(X\) is uniform on \(\{ 0,1\} ^{n}\), and \(Y\) independently flips its coordinates with probability \(\delta \).
There exists a universal real constant \(C\gt 0\) such that, for every \(n,k\in \mathbb N\) with \(k\geq 2\) and every \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) admitting a representation \(f(x)=\operatorname {sgn}_{+}(p((-1)^{x_1},\ldots ,(-1)^{x_n}))\) by a real multilinear polynomial \(p\) of degree at most \(k\),
Here \(\operatorname {Inf}(f)=\sum _{i=1}^{n}\Pr _X[f(X)\neq f(X^{\{ i\} })]\), \(X\) is uniform, \(X^{\{ i\} }\) flips coordinate \(i\), \(\log \) is the natural logarithm, real powers have positive bases, and \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\).
For every \(n,k\in \mathbb N\) with \(k\leq 1\), if \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) has a representation \(f(x)=\operatorname {sgn}_{+}(p((-1)^{x_1},\ldots ,(-1)^{x_n}))\) by a real multilinear polynomial \(p\) of degree at most \(k\), then there exist real \(c,a_1,\ldots ,a_n\) such that
where \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\).
There exists a universal real constant \(C\gt 0\) such that, for every \(n,k\in \mathbb N\), including \(n=0\) and \(k=0\), and every \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) represented as \(f(x)=\operatorname {sgn}_{+}(p((-1)^{x_1},\ldots ,(-1)^{x_n}))\) by a real multilinear polynomial \(p\) of degree at most \(k\),
Here \(\operatorname {Inf}(f)=\sum _{i=1}^{n}\Pr _X[f(X)\neq f(X^{\{ i\} })]\), \(X\) is uniform, \(X^{\{ i\} }\) flips coordinate \(i\), \(\log \) is the natural logarithm, real powers have positive bases, and \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\).
For \(k,d\in \mathbb N\) and real \(E,\delta \) with \(0\leq E\leq d\) and \(0\lt \delta \leq 1/2\), setting \(m=\lfloor 1/\delta \rfloor \) gives
where \(\log \) is the natural logarithm and the powers with real exponents have positive bases.
For \(n,k\in \mathbb N\), let a degree-at-most-\(k\) polynomial threshold function mean \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) of the form \(f(x)=\operatorname {sgn}_{+}(p((-1)^{x_1},\ldots ,(-1)^{x_n}))\) for a real multilinear polynomial \(p\) of degree at most \(k\), with \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\); there exists a universal \(C_0\gt 0\) such that every such function satisfies
In addition, for every fixed \(k\in \mathbb N\) with \(k\gt 0\), there exist a real \(C_k\geq 0\) and a natural number \(d_k\) such that every dimension \(n\in \mathbb N\), every degree-at-most-\(k\) polynomial threshold function \(f\) and every \(0\lt \delta \leq 1/2\) satisfy
Here \(\operatorname {Inf}(f)=\sum _{i=1}^{n}\Pr _X[f(X)\neq f(X^{\{ i\} })]\), \(X\) is uniform, \(X^{\{ i\} }\) flips coordinate \(i\), \(\operatorname {NS}_{\delta }(f)=\Pr [f(X)\neq f(Y)]\) when \(Y\) independently flips each coordinate of \(X\) with probability \(\delta \), and \(\log \) is the natural logarithm.
For every fixed \(k\in \mathbb N\), including \(k=0\), the degree-at-most-\(k\) polynomial threshold functions form a uniformly noise-stable family: there exists \(\varepsilon :\mathbb R\to \mathbb R\) with \(0\leq \varepsilon (\delta )\leq 1\) for \(0\leq \delta \leq 1/2\) and \(\lim _{\delta \to 0^{+}}\varepsilon (\delta )=0\) such that
for every \(n\in \mathbb N\), \(0\lt \delta \leq 1/2\) and \(f:\{ 0,1\} ^{n}\to \{ -1,1\} \) admitting a representation \(f(x)=\operatorname {sgn}_{+}(p((-1)^{x_1},\ldots ,(-1)^{x_n}))\) with a real multilinear polynomial \(p\) of degree at most \(k\). Here \(\operatorname {sgn}_{+}(t)=1\) for \(t\geq 0\) and \(-1\) for \(t\lt 0\), \(X\) is uniform on \(\{ 0,1\} ^{n}\) and \(Y\) independently flips each coordinate of \(X\) with probability \(\delta \).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). For every \(S\subseteq [n]\), \(\widehat P(S)=a_S\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). If \(k\ge 0\) is an integer and \(a_S=0\) whenever \(|S|\gt k\), then \(\widehat P(S)\ne 0\) implies \(|S|\le k\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). Then \(\mathbb E[P(x)P(x)]=\sum _{S\subseteq [n]}a_S^2\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). The second moment satisfies \(\mathbb E[P^2]=\sum _{S\subseteq [n]}a_S^2\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). For every real cube function \(u\), \(\mathbb E[Pu]=\sum _{S\subseteq [n]}a_S\widehat u(S)\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). For every \(S\subseteq [n]\), \(|a_S|\le \mathbb E|P|\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). If \(f(x)=1\) when \(P(x)\ge 0\) and \(f(x)=-1\) when \(P(x)\lt 0\), then \(\mathbb E|P|=\mathbb E[Pf]\).
Let \(n\ge 0\) and equip \(\{ 0,1\} ^n\) with its uniform probability measure; write \(\mathbb E\) for uniform expectation. For \(S\subseteq [n]\), put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for real cube functions \(h\). Let \(P(x)=\sum _{S\subseteq [n]}a_S\chi _S(x)\) with real coefficients \(a_S\). If \(a_S=0\) for \(|S|\gt 1\), then \(P(x)=a_\varnothing +\sum _{i=1}^n a_{\{ i\} }(-1)^{x_i}\) for every \(x\in \{ 0,1\} ^n\).
Let \(s\gt 0\) be an integer and \(Y_1,\ldots ,Y_s\) be independent real functions on a probability space \((\Omega ,\mu )\), each measurable up to a null set and valued in \([-1,1]\) almost surely, with common mean \(\mathbb E Y_i=m\). Write \(\Pr (A)=\mu ^*(A)\) for outer probability. For every real \(\varepsilon \ge 0\), \(\Pr [s\varepsilon \le \sum _iY_i-sm]\le \exp (-s\varepsilon ^2/2)\).
Let \(s\gt 0\) be an integer and \(Y_1,\ldots ,Y_s\) be independent real functions on a probability space \((\Omega ,\mu )\), each measurable up to a null set and valued in \([-1,1]\) almost surely, with common mean \(\mathbb E Y_i=m\). Write \(\Pr (A)=\mu ^*(A)\) for outer probability. For every real \(\varepsilon \ge 0\), \(\Pr [s\varepsilon \le |\sum _iY_i-sm|]\le 2\exp (-s\varepsilon ^2/2)\).
Let \((\Omega ,\mu )\) be a probability space, let \(J\) be a finite index set, and let \((A_j)_{j\in J}\) be arbitrary subsets of \(\Omega \). If a real number \(B\) satisfies \(\mu ^*(A_j)\le B\) for every \(j\) and \(|J|B\lt 1\), then some \(\omega \in \Omega \) belongs to none of the \(A_j\), where \(\mu ^*\) denotes the outer measure of \(\mu \).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). Write \(\operatorname {sgn}_+(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). For a real cube function \(f\) with \(L=\sum _{S\subseteq [n]}|\widehat f(S)|\gt 0\), there exists a probability distribution \(\pi \) on the subsets of \([n]\) such that \(\mathbb E_{S\sim \pi }[\operatorname {sgn}_+(\widehat f(S))\chi _S(x)]=f(x)/L\) for every cube point \(x\).
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). If a real cube function \(f\) is not identically zero, then its spectral one-norm \(\sum _{S\subseteq [n]}|\widehat f(S)|\) is strictly positive.
Let \(n\ge 0\) and give \(\{ 0,1\} ^n\) uniform probability; write \(\mathbb E\) for expectation. Put \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) and \(\widehat h(S)=\mathbb E[h\chi _S]\) for \(S\subseteq [n]\) and real cube functions \(h\). Write \(\operatorname {sgn}_+(t)=1\) for \(t\ge 0\) and \(-1\) for \(t\lt 0\). Let \(n\gt 0\), let \(f\) be a real cube function not identically zero, put \(L=\sum _S|\widehat f(S)|\), and let \(\delta \gt 0\) and an integer \(s\ge 0\) satisfy \(4nL^2/\delta ^2\le s\). There exists a sequence of subsets \(T_1,\ldots ,T_s\subseteq [n]\), with repetitions allowed, such that \(|f(x)-(L/s)\sum _{j=1}^s\operatorname {sgn}_+(\widehat f(T_j))\chi _{T_j}(x)|\lt \delta \) for every \(x\in \{ 0,1\} ^n\).
Fix natural numbers \(n\), \(d\), \(w\) and \(K\). If every \((n,d,w)\)-admissible code \(C\) satisfies \(\left\lvert C\right\rvert \le K\), then the maximum size of an \((n,d,w)\)-admissible code satisfies \(A_0(n,d,w) \le K\).
An \(\texttt{AdaptiveAdversary}\ N\) is an online strategy for the adversary: it provides a loss-selection function \(\mathrm{loss} : (t : \mathbb {N}) \to \texttt{LossHistory}\ N\ t \to (\mathrm{Fin}\, N \to \mathbb {R}) \to \mathrm{Fin}\, N \to \mathbb {R}\) that maps the current round index, the past loss history, and the learner’s current mixed strategy to a loss vector, subject to the validity condition that every coordinate of the chosen loss vector lies in \([0, 1]\).
Let \(D\) be a digraph on a finite vertex set \(V\), and let \(A\) be its \(0/1\) adjacency matrix over \(\mathbb {N}\), so \(A_{uv} = 1\) when there is an arc from \(u\) to \(v\) and \(A_{uv} = 0\) otherwise. Then for every \(k \in \mathbb {N}\) and all vertices \(u, v \in V\), the entry \((A^k)_{uv}\) equals the number of directed walks of length \(k\) from \(u\) to \(v\) in \(D\), that is, the number of sequences of \(k\) arrows \(u = w_0 \to w_1 \to \cdots \to w_k = v\).
Let \(G\) be a finite two-player zero-sum game with payoff matrix \(A\), having \(M \ge 2\) row actions and at least one column action, and let \(\varepsilon \gt 0\). Then there exist a mixed row strategy \(p\) and a mixed column strategy \(q\) that form an \(\varepsilon \)-approximate saddle point: for every pure row \(i\) and every pure column \(j\),
That is, up to an additive \(\varepsilon \), the expected payoff of \(p\) against any pure column is at least the expected payoff of any pure row against \(q\).
A social welfare function \(f\) is acyclic if no profile of transitive voter orderings produces a Condorcet cycle in society’s preferences. A cycle occurs when \(f(\mathrm{abVotes}(p)) = f(\mathrm{bcVotes}(p)) = f(\mathrm{caVotes}(p)) = 1\) (the cycle \(a\gt b\gt c\gt a\)) or all equal \(-1\) (the reverse cycle).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function that is odd, \(\pm 1\)-valued, and acyclic. Then its Fourier correlation function satisfies
Let \(n\) be a natural number and let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a social welfare function that aggregates the pairwise opinions of \(n\) voters. Suppose \(f\) is odd, meaning \(f(\bar x) = -f(x)\) for every \(x\), where \(\bar x\) is the bitwise complement of \(x\); is \(\pm 1\)-valued, meaning \(f(x)\in \{ -1,1\} \) for every \(x\); is unanimous, meaning \(f\) takes the value \(1\) at the all-zeros input; and is acyclic, meaning that no profile of voter orderings produces a Condorcet cycle. Here a profile assigns to each of the \(n\) voters one of the six linear orderings of three alternatives \(a,b,c\), and from it one forms three vote vectors in \(\{ 0,1\} ^n\) whose \(i\)-th entries record voter \(i\)’s preference in the \(a\)-vs-\(b\), \(b\)-vs-\(c\), and \(c\)-vs-\(a\) comparisons; acyclicity requires that \(f\) never sends all three of these vectors to \(1\), and never sends all three to \(-1\). Then \(f\) is a dictatorship: there is a voter \(i\) such that \(f(x) = (-1)^{x_i}\) for all \(x\).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a \(\pm 1\)-valued Boolean function that is odd, meaning \(f(\lnot x)=-f(x)\) for every \(x\), and suppose its Fourier correlation function attains the value
Then \(\hat f(S)=0\) for every subset \(S\subseteq [n]\) with \(|S|\neq 1\); that is, all Fourier–Walsh weight of \(f\) is carried by the singleton frequencies.
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function that is odd, so that \(f(\lnot x)=-f(x)\) for all \(x\), and \(\pm 1\)-valued, so that \(f(x)\in \{ -1,1\} \) for all \(x\). Then its Fourier correlation function satisfies
where \(\hat f(S)\) denotes the Fourier–Walsh coefficient of \(f\) at frequency \(S\).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function that is odd, is \(\pm 1\)-valued, and is unanimous. Suppose further that \(f\) has degree one, in the sense that its Fourier–Walsh coefficient \(\hat f(S)\) vanishes for every set \(S\subseteq [n]\) with \(\left\lvert S\right\rvert \neq 1\). Then \(f\) is a dictatorship: there is a coordinate \(i\) with \(f=\mathrm{dict}_i\).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function. For a profile \(p\) that assigns each of the \(n\) voters one of the six strict orderings of the three alternatives \(a,b,c\), let \(x^{ab}(p)\in \{ 0,1\} ^n\) record each voter’s preference in the \(a\)-vs-\(b\) comparison and let \(x^{ca}(p)\in \{ 0,1\} ^n\) record each voter’s preference in the \(c\)-vs-\(a\) comparison. Averaging uniformly over all \(6^n\) profiles, the expected product of \(f\) at these two vote vectors equals the Fourier correlation function of \(f\):
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function, and for each profile \(p:\mathrm{Fin}\, n\to \{ 0,\dots ,5\} \) write \(x_{bc}(p),\, x_{ca}(p)\in \{ 0,1\} ^n\) for the \(b\)-vs-\(c\) and \(c\)-vs-\(a\) vote vectors it induces. Then the average of the product \(f\bigl(x_{bc}(p)\bigr)\, f\bigl(x_{ca}(p)\bigr)\) over all \(6^n\) profiles equals the Fourier correlation function of \(f\):
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function. Suppose each of the \(n\) voters chooses, independently and uniformly, one of the six strict orderings of three candidates \(a,b,c\); for such a profile \(p\) write \(x_p\in \{ 0,1\} ^n\) for the vector whose \(i\)-th entry records voter \(i\)’s preference in the \(a\)-versus-\(b\) comparison, and \(y_p\in \{ 0,1\} ^n\) for the vector recording each voter’s preference in the \(b\)-versus-\(c\) comparison. Then, averaging the product of the values of \(f\) on these two vectors over all \(6^n\) profiles,
where \(\hat f(S)\) denotes the Fourier–Walsh coefficient of \(f\) at frequency \(S\).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function, and let \(\mathrm{votes}_1,\mathrm{votes}_2\) each assign to every profile \(p\in (\mathrm{Fin}\, 6)^{\, n}\) a point of the Boolean hypercube \(\{ 0,1\} ^n\). Suppose the pair satisfies the kernel identity
for all \(S,T\subseteq [n]\), where the sum runs over all \(6^n\) profiles. Then
Let \(A \subseteq \mathrm{Fin}\, n\) be a finite set of indices and let \(g : \mathrm{Fin}\, n \to \mathbb {R}\). Then
so that a product taken only over the members of \(A\) may equivalently be written as a product over all of \(\mathrm{Fin}\, n\) in which each factor outside \(A\) is replaced by \(1\).
Fix \(n\), and for each profile \(p \in \{ 0,\dots ,5\} ^n\) let \(v_{ab}(p), v_{bc}(p) \in \{ 0,1\} ^n\) be the vote vectors recording, voter by voter, each voter’s pairwise preference in the \(a\)-vs-\(b\) and \(b\)-vs-\(c\) comparisons, and for \(S \subseteq [n]\) let \(\chi _S\) be the associated Walsh–Fourier character. Then for all subsets \(S, T \subseteq [n]\), the uniform average of the product \(\chi _S(v_{ab}(p))\, \chi _T(v_{bc}(p))\) over the \(6^n\) profiles satisfies
Fix a natural number \(n\) and consider profiles \(p\) of \(n\) voters, each of whom reports one of the six strict orderings of the three alternatives \(a, b, c\); there are \(6^n\) such profiles. For a profile \(p\), let \(x^{ab}(p) \in \{ 0,1\} ^n\) be the vector recording each voter’s preference in the \(a\)-vs-\(b\) comparison, and \(x^{ca}(p) \in \{ 0,1\} ^n\) the vector recording each voter’s preference in the \(c\)-vs-\(a\) comparison. Then for any two subsets \(S, T \subseteq \{ 1,\dots ,n\} \), the averaged product of the corresponding Walsh–Fourier characters satisfies
where the sum ranges over all \(6^n\) profiles.
Let \(S, T \subseteq [n]\), and for a profile \(p \colon \mathrm{Fin}\, n \to \mathrm{Fin}\, 6\) write \(x(p)\) for its \(bc\)-vote vector and \(y(p)\) for its \(ca\)-vote vector, each lying in \(\{ 0,1\} ^n\). Averaging the product of Walsh characters over all \(6^n\) profiles gives
Let \(x, y : \mathrm{Fin}\, 6 \to \mathrm{Bool}\) be two preference assignments whose sign encodings have balanced marginals, \(\sum _{k=1}^{6}\sigma (x_k) = \sum _{k=1}^{6}\sigma (y_k) = 0\), and whose signed correlation is \(\sum _{k=1}^{6}\sigma (x_k)\, \sigma (y_k) = -2\). For a profile \(p : \mathrm{Fin}\, n \to \mathrm{Fin}\, 6\), write \(x\circ p\) and \(y\circ p\) for the points of the Boolean hypercube whose \(i\)-th coordinates are \(x_{p(i)}\) and \(y_{p(i)}\). Then for all subsets \(S, T \subseteq [n]\), the average of the product of Walsh characters over all \(6^n\) profiles satisfies
Label the six strict orderings of three alternatives \(a,b,c\) by \(k\in \{ 0,\dots ,5\} \). For each ordering let \(x_k\in \{ -1,+1\} \) record the \(a\)-versus-\(b\) comparison, taking the value \(+1\) when \(a\) is preferred to \(b\) and \(-1\) when \(b\) is preferred to \(a\), and let \(y_k\in \{ -1,+1\} \) record the \(b\)-versus-\(c\) comparison in the same way, with \(+1\) when \(b\) is preferred to \(c\) and \(-1\) otherwise. Then
For each of the six strict orderings \(k \in \{ 0,\dots ,5\} \) of three alternatives \(a,b,c\), let \(\sigma \) be the sign encoding, sending \(\mathrm{false}\) to \(+1\) and \(\mathrm{true}\) to \(-1\). Write \(y_k = \sigma (p_k)\), where \(p_k\) is the Boolean preference of ordering \(k\) in the \(b\)-vs-\(c\) comparison, and \(z_k = \sigma (q_k)\), where \(q_k\) is the preference in the \(c\)-vs-\(a\) comparison. Then
For each of the six strict orderings \(k \in \{ 0,\dots ,5\} \) of three alternatives \(a,b,c\), write \(p_{ab}(k)\) and \(p_{ca}(k)\) for the Boolean preferences of ordering \(k\) in the \(a\)-vs-\(b\) and \(c\)-vs-\(a\) comparisons, and let \(\sigma \) be the sign encoding sending \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\). Then the sum over all six orderings of the products of the encoded signs is
There are six strict orderings \(k \in \{ 0,1,\dots ,5\} \) of the three alternatives \(a,b,c\). For each such ordering, let \(p_{ab}(k)\) be its \(a\)-vs-\(b\) preference, the Boolean that records whether \(b\) is ranked above \(a\), and apply the sign encoding \(\sigma \), so that \(\sigma (p_{ab}(k))\) equals \(+1\) when ordering \(k\) ranks \(a\) above \(b\) and \(-1\) when it ranks \(b\) above \(a\). Then these six signed preferences sum to zero:
Consider the six strict orderings of three alternatives \(a\), \(b\), and \(c\), indexed by \(k \in \{ 0,\dots ,5\} \), and let \(p(k)\) denote the \(b\)-versus-\(c\) pairwise preference of ordering \(k\). Under the sign encoding \(\sigma \), these preferences sum to zero across all six orderings:
For each of the six strict orderings \(k\in \{ 0,1,\dots ,5\} \) of three alternatives, let \(b_k\in \mathrm{Bool}\) record the \(c\)-vs-\(a\) preference of ordering \(k\), and encode it as a sign by \(\sigma (b_k)=1\) when \(b_k\) is false and \(\sigma (b_k)=-1\) when \(b_k\) is true. Then these six signed preferences sum to zero:
Let \(G\) be a finite simple graph on a vertex set with at least two vertices, let \(D(G)\) be its associated digraph (each edge of \(G\) replaced by the two oppositely oriented arcs joining its ends), and let \(k\) be a natural number. Then \(D(G)\) is \(k\)-arc-connected if and only if \(k \le \kappa '(G)\), where \(\kappa '(G)\) denotes the edge connectivity of \(G\).
Given \(T\) distributions \(p_0, \dots , p_{T-1}\) over \(\mathrm{Fin}\, n\) (each presented as non-negative weights summing to \(1\)) and \(T \gt 0\), the average strategy is the mixed strategy with weights
Let \(N \ge 1\) and \(T \ge 1\) be integers, and let \(f(t,j)\) be a real number for each \(t \in \{ 1,\dots ,T\} \) and each \(j \in \{ 1,\dots ,N\} \). Then averaging over \(t\) before minimising over \(j\) can only increase the value:
Given a \(k \times d\) matrix \(A\) and a finite set \(V \subseteq \mathbb {R}^d\), BadPair \(\varepsilon \; V\; A\) holds when there exists some ordered pair \((u, v) \in V \times V\) whose images under \(A\) fail the distortion predicate, i.e. the projection distorts that pair by more than a factor of \(\varepsilon \).
For a parameter \(\varepsilon \in \mathbb {R}\), a matrix \(A\), and a vector \(x\in \mathbb {R}^d\), the predicate \(\texttt{BadSingle}\, \varepsilon \, A\, x\) holds when the squared-norm distortion is strictly too large:
This is the “bad event” whose probability must be controlled by the JL concentration argument.
Let \(L \in [0,1]\) and let \(\eta \) be a real number. Then the moment generating function of the shifted Bernoulli weight \(1 - L + L\, e^{-\eta }\) satisfies
Let \(L, x \in \mathbb {R}\) and suppose \(1 - L + L\, e^{-x} \gt 0\). Then the auxiliary function
is differentiable at \(x\) with derivative
This is the \(\varphi '\) computation used in the proof of bernoulli_mgf_bound.
Let \(L, x \in \mathbb {R}\) and suppose \(1 - L + L\, e^{-x} \gt 0\). Then the function
(the derivative from bernoulli_mgf_bound_step1) is differentiable at \(x\) with derivative
This is the \(\varphi ''\) computation used in the proof of bernoulli_mgf_bound.
Fix \(n\) and a real parameter \(p\) with \(0 \le p \le 1\), and let \(\mathrm{event}\) be any (decidable) property of restrictions on \(n\) variables. Then the probability that \(\mathrm{event}\) holds under a Bernoulli(\(p\)) random restriction equals its average over the fixed-size restriction models, weighted by the binomial distribution on the number of free variables:
where \(R_k(\mathrm{event})\) is the fraction of restrictions with exactly \(k\) free variables that satisfy \(\mathrm{event}\).
Fix a parameter \(p\) with \(0 \lt p \le 1\), positive integers \(w, s \gt 0\), and a tolerance \(\varepsilon \gt 0\). Then there exists \(N \in \mathbb {N}\) such that for every \(n \ge N\) with \(n \gt 0\) and every decidable event on restrictions of \(n\) variables, the following holds: if the fixed-size restriction probability satisfies
then the probability of the event under a Bernoulli(\(p\)) random restriction satisfies
Let \(n \ge 1\), let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le 1\), and let \(w, s\) be positive integers. Let \(\mathrm{event}\) be a decidable predicate on restrictions of \(n\) variables, and suppose that for every \(k \le n\) its probability under the fixed-size restriction model \(R_k\) satisfies
Then the probability of \(\mathrm{event}\) under a Bernoulli(\(p\)) random restriction satisfies
Let \(n\) be a natural number and let \(p\) be a real number with \(0 \le p \le 1\). Then the binomial probabilities \(\binom {n}{k}\, p^k\, (1-p)^{n-k}\) sum to one over all outcomes,
Let \(n\) be a natural number and let \(p\) be a real number with \(0 \lt p \le 1\). Summing the binomial probability mass function \(\binom {n}{k}\, p^k\, (1-p)^{n-k}\) over those integers \(k \in \{ 0, 1, \dots , n\} \) satisfying \(k \gt 2np\) yields at most \(e^{-np/3}\); that is, the probability that \(\mathrm{Bin}(n,p)\) exceeds twice its mean is bounded by \(e^{-np/3}\).
The probability of an event under the fixed-size restriction model \(R_k\), which is uniform over restrictions with exactly \(k\) free variables. It is defined as the fraction of restrictions with exactly \(k\) free variables that satisfy the event.
Let \(E\) be any event, that is, a property of restrictions on \(n\) variables, and let \(k\) be a natural number. Under the fixed-size restriction model \(R_k\) — the uniform distribution over those restrictions on \(n\) variables whose set of free variables has cardinality exactly \(k\) — the probability of \(E\) satisfies \(R_k(E) \le 1\).
Fix natural numbers \(n\) and \(k\), and let \(E\) be any event on restrictions of \(n\) variables. Then the probability of \(E\) under the fixed-size restriction model \(R_k\) — the uniform distribution over restrictions with exactly \(k\) free variables — is nonnegative:
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N \ge 1\) column actions, given by a payoff matrix \(A_{ij} \in [0,1]\), and let \(p\) be a mixed strategy for the row player, with weights \(p_i\). Write \(v(j) = \sum _i p_i A_{ij}\) for the expected payoff of \(p\) against the pure column \(j\), and let \(j^*\) be a column that minimises \(v\) over all \(N\) columns. Then \(v(j^*) \le v(j)\) for every column \(j\); that is, \(j^*\) is indeed a minimiser of the expected payoff against \(p\).
Let \(G\) be a finite two-player zero-sum game with a nonempty set of \(M\) row actions and \(N\) column actions, given by a payoff matrix \(A\) with \(0 \le A_{ij} \le 1\), and let \(q\) be a mixed strategy for the column player. For a pure row \(i\), write \(u(i) = \sum _{j} A_{ij}\, q_j\) for its expected payoff against \(q\), and let \(i^\ast \) be a row maximising \(u\). Then \(u(i) \le u(i^\ast )\) for every row \(i\).
Fix integers \(n, d, w\) with \(n \ge 1\), \(1 \le d\), and \(2d \le n\), and suppose that \(w \le J_2(n,d) = \tfrac {1}{2}\bigl(n - \sqrt{n\, (n-2d)}\bigr)\). Then the maximum size \(A_0(n,d,w)\) of an \((n,d,w)\)-admissible binary code — a set of binary words of length \(n\) in which any two distinct words are at Hamming distance at least \(d\) and every word has weight at most \(w\) — satisfies
Let \(n\), \(\ell \), and \(q\) be natural numbers with \(n \ge 1\) and \(q \ge 1\), and suppose that \(q \le \ell \), \(2\ell \le n\), and \(2q \le n\). Then
where the binomial coefficients are evaluated in \(\mathbb {R}\).
Let \(n, \ell , q\) be natural numbers with \(n \gt 0\), \(q \gt 0\), \(q \le \ell \), \(2\ell \le n\), and \(2q \le n\). Then the ratio of real-valued binomial coefficients satisfies
Fix a real number \(\tau \) with \(0 \le \tau \le \tfrac 12\). Then the normalized base-\(2\) logarithm of the lower binomial tail sum converges to the binary entropy of \(\tau \):
The BLR acceptance probability of \(f : \{ 0,1\} ^n \to \{ 0,1\} \) is
where \(x, y\) are drawn uniformly from \(\{ 0,1\} ^n\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function, and write \(F(x) = (-1)^{f(x)} \in \{ \pm 1\} \) for its sign lift. Then the BLR acceptance probability of \(f\), the probability that \(F(x \oplus y) = F(x)\, F(y)\) when \(x\) and \(y\) are drawn independently and uniformly from \(\{ 0,1\} ^n\), satisfies
where \(x \oplus y\) denotes the componentwise XOR.
Fix \(n \in \mathbb {N}\) and a Boolean-valued function \(f : \{ 0,1\} ^n \to \{ 0,1\} \), and let \(g : \{ 0,1\} ^n \to \mathbb {R}\) be its \(\pm 1\) lift, \(g(x) = (-1)^{f(x)}\). Then for all \(x, y \in \{ 0,1\} ^n\),
where \(x \oplus y\) denotes the componentwise XOR and the left-hand side is \(1\) when the equality holds and \(0\) otherwise.
Let \(n\) be a natural number and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be linear, meaning that \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\). Then the BLR acceptance probability of \(f\) equals \(1\); that is, when \(x\) and \(y\) are drawn uniformly and independently from \(\{ 0,1\} ^n\), the test always accepts.
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a function that respects componentwise XOR, meaning \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), and let \(g : \{ 0,1\} ^n \to \mathbb {R}\) be its \(\pm 1\) lift, \(g(x) = (-1)^{f(x)}\). Then \(g\) turns XOR into multiplication:
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a linear function, meaning that \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), where \(\oplus \) denotes componentwise XOR. Then for all \(x, y \in \{ 0,1\} ^n\) we have \(f(x \oplus y) = f(x) \oplus f(y)\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) and let \(\varepsilon \in \mathbb {R}\) satisfy \(0 \le \varepsilon \le 1\). Suppose \(f\) is \(\varepsilon \)-far from linear, meaning that every linear function \(g : \{ 0,1\} ^n \to \{ 0,1\} \) (one satisfying \(g(x \oplus y) = g(x) \oplus g(y)\) for all \(x, y\)) has distance \(\Pr _{x}[f(x) \ne g(x)] \ge \varepsilon \) from \(f\). Then the BLR acceptance probability of \(f\) satisfies
where \(x\) and \(y\) range uniformly over \(\{ 0,1\} ^n\); equivalently, the test rejects \(f\) with probability at least \(\varepsilon \).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function and let \(\varepsilon \in \mathbb {R}\) with \(0 \le \varepsilon \le 1\), and suppose \(f\) is \(\varepsilon \)-far from linear: for every linear function \(g : \{ 0,1\} ^n \to \{ 0,1\} \) one has \(\operatorname{dist}(f,g) \ge \varepsilon \). Writing \(F = (-1)^{f}\) for the \(\pm 1\) lift of \(f\) and \(\hat F(S) = \langle F, \chi _S\rangle \) for its Fourier coefficient at \(S \subseteq [n]\), the sum of the cubed coefficients over all subsets satisfies
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be \(\varepsilon \)-far from linear, meaning that \(0 \le \varepsilon \le 1\) and every linear function \(g : \{ 0,1\} ^n \to \{ 0,1\} \) (one satisfying \(g(x \oplus y) = g(x) \oplus g(y)\) for all \(x,y\)) has distance \(\Pr _{x}[f(x) \ne g(x)] \ge \varepsilon \) from \(f\). Write \(\hat F(S) = \mathbb {E}[F \cdot \chi _S]\) for the Fourier coefficient of \(F = (-1)^{f}\), the \(\pm 1\) lift of \(f\), at a subset \(S \subseteq [n]\). Then for every such \(S\),
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function, and let \(g : \{ 0,1\} ^n \to \mathbb {R}\) be its \(\pm 1\) lift, given by \(g(x) = (-1)^{f(x)}\). Then the Fourier coefficients \(\hat g(S) = \langle g, \chi _S\rangle \), taken over all subsets \(S \subseteq [n]\), satisfy
Let \(f,g:\{ 0,1\} ^n\to \{ 0,1\} \) be Boolean-valued functions, let \(S\subseteq [n]\), and let \(\varepsilon \in \mathbb {R}\). Write \(f^{\pm },g^{\pm }:\{ 0,1\} ^n\to \mathbb {R}\) for their \(\pm 1\) lifts, given by \(x\mapsto (-1)^{f(x)}\) and \(x\mapsto (-1)^{g(x)}\). Suppose that \(g^{\pm }\) coincides with the Walsh character \(\chi _S\), and that \(\operatorname{dist}(f,g)=\Pr _{x}[f(x)\ne g(x)]\ge \varepsilon \). Then the Fourier coefficient of \(f^{\pm }\) at \(S\) satisfies
Let \(f, g : \{ 0,1\} ^n \to \{ 0,1\} \) be Boolean-valued functions on the hypercube, and write \(F, G : \{ 0,1\} ^n \to \mathbb {R}\) for their \(\pm 1\) lifts, given by \(F(x) = (-1)^{f(x)}\) and \(G(x) = (-1)^{g(x)}\). Then for every point \(x \in \{ 0,1\} ^n\),
so the product equals \(+1\) where \(f\) and \(g\) agree at \(x\) and \(-1\) where they disagree.
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be \(\varepsilon \)-far from linear, meaning that \(0 \le \varepsilon \le 1\) and every linear function \(g : \{ 0,1\} ^n \to \{ 0,1\} \) (one satisfying \(g(x \oplus y) = g(x) \oplus g(y)\) for all \(x,y\)) has \(\Pr _{x}[f(x) \ne g(x)] \ge \varepsilon \) under the uniform distribution. Let \(F(x) = (-1)^{f(x)}\) be the \(\pm 1\) lift of \(f\). Then for every subset \(S \subseteq [n]\), the Fourier coefficient of \(F\) at \(S\) satisfies
where \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\) is the Walsh character indexed by \(S\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function. Then \(f\) is linear — that is, \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), where \(\oplus \) denotes bitwise XOR — if and only if there exists a set \(S \subseteq [n]\) for which the \(\pm 1\) lift \(x \mapsto (-1)^{f(x)}\) coincides with the Walsh character \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\).
Fix \(n\) and a subset \(S \subseteq [n]\), and let \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\) be the associated Walsh character on the hypercube \(\{ 0,1\} ^n\). Then for all \(x, y \in \{ 0,1\} ^n\),
where \(x \oplus y\) denotes the componentwise XOR.
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function whose \(\pm 1\) lift coincides with a Walsh character, say \(x \mapsto (-1)^{f(x)}\) equals \(\chi _S\) for some \(S \subseteq [n]\), and fix points \(x, y \in \{ 0,1\} ^n\). If the character is multiplicative at this pair, meaning \(\chi _S(x \oplus y) = \chi _S(x)\, \chi _S(y)\) where \(x \oplus y\) denotes the componentwise XOR, then
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function whose \(\pm 1\) lift \(x \mapsto (-1)^{f(x)}\) coincides with the Walsh character \(\chi _S\) for some subset \(S \subseteq [n]\). Then for all \(x, y \in \{ 0,1\} ^n\),
where \(x \oplus y\) denotes the componentwise XOR of \(x\) and \(y\) and \(f(x) \oplus f(y)\) the XOR of the two Boolean outputs.
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a linear Boolean function, meaning \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), where \(\oplus \) is componentwise XOR. For each coordinate \(i\), let \(e_i \in \{ 0,1\} ^n\) denote the standard basis vector that is \(1\) in coordinate \(i\) and \(0\) elsewhere. Then for every \(x \in \{ 0,1\} ^n\), the value \(f(x)\) is \(1\) if the number of coordinates \(i\) with \(x_i = 1\) and \(f(e_i) = 1\) is odd, and \(0\) if that number is even; equivalently, \(f(x)\) is the parity of \(\lvert \{ \, i : x_i = 1 \text{ and } f(e_i) = 1 \, \} \rvert \).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a linear Boolean function, meaning \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), where \(\oplus \) is componentwise XOR. For a subset \(s \subseteq \{ 0, \dots , n-1\} \), let \(\mathbf{1}_s \in \{ 0,1\} ^n\) be its indicator vector, whose \(i\)-th coordinate is \(1\) when \(i \in s\) and \(0\) otherwise, and for each index \(i\) let \(e_i \in \{ 0,1\} ^n\) be the standard basis vector, whose \(i\)-th coordinate is \(1\) and all others \(0\). Then for every \(s\) the value \(f(\mathbf{1}_s)\) is \(0\) when
and \(1\) otherwise; that is, \(f(\mathbf{1}_s)\) equals the parity of the number of \(i \in s\) with \(f(e_i) = 1\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a linear Boolean function, meaning \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), and for each coordinate \(i\) let \(e_i \in \{ 0,1\} ^n\) denote the standard basis vector that is \(1\) in position \(i\) and \(0\) elsewhere. Then for every \(x \in \{ 0,1\} ^n\), the value \(f(x)\) equals the parity of the number of coordinates \(i\) with \(x_i = 1\) and \(f(e_i) = 1\); that is, \(f(x) = 1\) if that count is odd and \(f(x) = 0\) if it is even.
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function on the hypercube, write \(x \oplus y\) for the componentwise XOR of \(x, y \in \{ 0,1\} ^n\), and let \(\mathbb {E}\) denote expectation over a uniformly random point of \(\{ 0,1\} ^n\). Then
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function on the hypercube, and let \(g = (-1)^{f}\) denote its \(\pm 1\) lift. Then the double uniform expectation of the triple product of \(g\) over independent inputs and their componentwise XOR equals the sum of the cubes of the Fourier coefficients of \(g\):
where \(x, y\) range uniformly over \(\{ 0,1\} ^n\), \((x \oplus y)_i = x_i \oplus y_i\), and \(\hat{g}(S) = \langle g, \chi _S \rangle \) is the Fourier coefficient at \(S \subseteq [n]\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function, and let \(g : \{ 0,1\} ^n \to \mathbb {R}\) be its \(\pm 1\) lift, \(g(x) = (-1)^{f(x)}\). Writing \(g * g\) for the self-convolution \((g * g)(x) = \mathbb {E}_{y}\bigl[g(y)\, g(x \oplus y)\bigr]\) with respect to the uniform measure, and \(\hat g(S) = \mathbb {E}[g\, \chi _S]\) for the Fourier coefficient of \(g\) at \(S\), one has for every \(x \in \{ 0,1\} ^n\)
where \(\chi _S(x) = \prod _{i \in S} (-1)^{x_i}\) is the Walsh character of \(S\). That is, the self-convolution of \(g\) has Fourier coefficients \(\hat g(S)^2\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function, and let \(g\) denote its \(\pm 1\) lift, so that \(g(x) = (-1)^{f(x)}\). Suppose that the convolution of \(g\) with itself admits the Fourier expansion
where \(\chi _S\) is the Walsh character of \(S\) and \(\hat{g}(S) = \langle g, \chi _S\rangle \). Then
the expectation being taken uniformly over \(x \in \{ 0,1\} ^n\).
Let \(f\colon A \to \mathbb {N}\) and \(g\colon B \to \mathbb {N}\) be functions into the natural numbers, let \(h\colon A \to B\) be a map, and let \(c_1,\dots ,c_m\) be a finite list of elements of \(A\). If \(g(h(c)) = f(c)\) for every entry \(c\) of the list, then
where each sum is taken over the list in order with initial value \(0\).
Let \(f\colon \alpha \to \mathbb {N}\), \(g\colon \beta \to \mathbb {N}\), and \(h\colon \alpha \to \beta \) be functions into the natural numbers, let \(cs\) be a finite list of elements of \(\alpha \), and let \(k\in \mathbb {N}\). If \(g(h(c)) \le k\, f(c)\) holds for every entry \(c\) of \(cs\), then, summing over the entries of \(cs\) (with multiplicity),
Fix a number of variables \(n\) and a Boolean assignment \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\). For any finite list of circuit literals \(\ell _1, \ldots , \ell _k\) on \(n\) variables, the right fold of logical conjunction over their evaluations, taken with initial value \(\mathrm{true}\) — that is, \(\ell _1.\mathrm{eval}(x) \wedge \bigl(\ell _2.\mathrm{eval}(x) \wedge \cdots \wedge (\ell _k.\mathrm{eval}(x) \wedge \mathrm{true})\bigr)\) — equals the term evaluation, on \(x\), of the list obtained by converting each circuit literal to its corresponding switching-lemma literal.
Let \(\alpha \) and \(\beta \) be types, let \(f\colon \alpha \to \{ \mathrm{true},\mathrm{false}\} \) and \(g\colon \beta \to \{ \mathrm{true},\mathrm{false}\} \) be Boolean-valued functions, let \(h\colon \alpha \to \beta \), and let \(cs\) be a finite list of elements of \(\alpha \). Suppose that \(g(h(c)) = f(c)\) for every \(c\) appearing in \(cs\). Then the Boolean conjunction of \(g\) over the mapped list, \(\bigwedge _{c\in cs} g(h(c))\), equals the Boolean conjunction of \(f\) over \(cs\), namely \(\bigwedge _{c\in cs} f(c)\), where in each case the conjunction is folded from the right starting from \(\mathrm{true}\).
Let \(\ell _1, \dots , \ell _k\) be a list of circuit literals on \(n\) variables and let \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\) be an assignment. Then the right fold of disjunction over the literal evaluations, starting from \(\mathrm{false}\),
equals the clause evaluation on \(x\) of the term obtained by converting each circuit literal to its switching-lemma literal; that is, both equal \(\mathrm{true}\) precisely when some literal in the list holds under \(x\).
Let \(f : \alpha \to \{ \text{true},\text{false}\} \) and \(g : \beta \to \{ \text{true},\text{false}\} \) be Boolean-valued functions, let \(h : \alpha \to \beta \), and let \(c_1, \dots , c_n\) be a finite list of elements of \(\alpha \). If \(g(h(c)) = f(c)\) for every element \(c\) of the list, then the disjunction \(g(h(c_1)) \vee \cdots \vee g(h(c_n))\) of the values of \(g\) over the mapped list equals the disjunction \(f(c_1) \vee \cdots \vee f(c_n)\) of the values of \(f\) over the original list, where each empty disjunction is taken to be \(\text{false}\).
A gate over alphabet \(\alpha \) with input domain domain consists of a gate operation op (an arity type \(\iota \) together with a function \((\iota \to \alpha ) \to \alpha \)) and a wiring map \(\iota \to \texttt{domain}\) that assigns to each input slot the node it reads from.
A single operation in a feedforward circuit over an alphabet \(\alpha \): it consists of an arity type \(\iota \) together with a function \((\iota \to \alpha ) \to \alpha \) that computes the gate’s output value from the values supplied on its input slots.
A GatePolyFamily for a gate operation \(\mathrm{op}\) packages a nonempty finite seed type together with, for each tuple of incoming polynomials and each seed, an approximating polynomial in \(n\) variables over \(\mathbb {Z}/p\) subject to two guarantees: its total degree is at most \((p-1)\ell \cdot \sup _i \deg (\text{polys}_i)\), and for every Boolean input \(x\) at which all incoming polynomials evaluate into \(\{ 0,1\} \), the number of seeds on which the approximator disagrees with the true gate value, multiplied by \(2^{\ell }\), is at most the number of seeds. Crucially the seed type depends only on the gate, not on the incoming polynomials.
A feedforward circuit from inp to out over \(\alpha \) consists of a depth \(d \in \mathbb {N}\), a family of node types indexed by the layers \(\mathrm{Fin}\, (d+1)\), and, for every layer \(i \lt d\), a gate for each node of layer \(i+1\) whose inputs are wired to nodes of layer \(i\). Layer \(0\) is required to be the input type inp and the last layer the output type out.
For an input assignment \(xs : \texttt{inp} \to \alpha \), the value of a node at layer \(d\) is defined by recursion on \(d\): a layer-\(0\) node is read off directly from \(xs\), and a node at layer \(n+1\) evaluates its gate on the values of the layer-\(n\) nodes.
Given a gate \(g\) and an assignment \(xs\) of values to the nodes of its input domain, the value of \(g\) is \(g.\texttt{op}.\texttt{func}\, (xs \circ g.\texttt{inputs})\), i.e. the gate operation applied to the values read along the input wires.
Given a Boolean feedforward circuit \(F\), a flag \(\texttt{isAnd}\) on each non-input node and a finiteness witness \(\texttt{gfin}\) for each gate’s arity, \(F.\texttt{IsAndOrGate}\) asserts that every gate operation is the conjunction of its inputs when the flag is true and their disjunction when it is false, the inputs being enumerated by \(\texttt{gfin}\). This is the gate restriction under which \(F\) can be converted into a BoolCircuit.TreeCircuit.
Let \(F\) be a Boolean feedforward circuit with input variables indexed by \(\mathrm{Fin}\, n\) and output nodes of type \(\mathrm{out}\), equipped with a flag assigning to each non-input node the label AND or OR and a witness that each gate has finite arity. Suppose \(F\) is an AND/OR circuit: every gate computes the conjunction of the values on its input wires when its node is labelled AND, and their disjunction when it is labelled OR. Then for every output node \(o\) and every assignment \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\), the tree-shaped Boolean circuit obtained by unrolling \(o\) from the last layer down to the inputs, evaluated at \(x\), agrees with the value that \(F\) itself assigns to \(o\) under the input assignment \(x\).
Let \(F\) be a layered feedforward circuit over the Boolean alphabet with input type \(\mathrm{Fin}\, n\) and output type \(\mathrm{out}\), of depth \(d\), in which each gate is designated an AND gate or an OR gate. Suppose every gate has finite input arity, and that there is a bound \(k \in \mathbb {N}\) such that every gate reads from at most \(k\) input wires. Then for each output node \(o\), the tree-unrolled Boolean circuit obtained from \(o\) has size (its total number of AND/OR gates and literal leaves) at most \((k+1)^{d}\).
Let \(\ell _1,\dots ,\ell _m\) be a list of circuit literals over \(n\) variables whose variable indices are pairwise distinct. Then any two literals \(l_1\) and \(l_2\) in this list that have the same variable index are equal; that is, if \(l_1\) and \(l_2\) both occur in the list and \(l_1.\mathrm{idx} = l_2.\mathrm{idx}\), then \(l_1 = l_2\).
For a literal \(l\) on \(n\) variables and an assignment \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\), the value \(l.\mathrm{eval}\, x\) is \(x_{l.\mathrm{idx}}\) when \(l\) is positive (\(l.\mathrm{sign} = \mathrm{true}\)) and \(\lnot x_{l.\mathrm{idx}}\) when \(l\) is negated.
Let \(l\) be a circuit literal on \(n\) variables and let \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\) be an assignment of Boolean values to those variables. Then evaluating \(l\) as a circuit literal at \(x\) yields the same Boolean value as first converting \(l\) to the corresponding switching-lemma literal and evaluating that at \(x\).
A normal-form AND circuit on \(n\) variables: either a base clause given by a list of literals whose variable indices are pairwise distinct (a Nodup proof on \(lits.\mathrm{map}\; \mathrm{Lit.idx}\)), or a node that ANDs together a list of normal-form OR circuits. Together with \(\mathrm{NOrCircuit}\) this forms a strictly alternating AND/OR normal form.
Fix \(n\), and let \(lits\) be a list of literals over the variables \(\{ 0,\dots ,n-1\} \) whose underlying variable indices, obtained by taking the index of each literal in \(lits\), are pairwise distinct. If a normal-form AND circuit on \(n\) variables equals the base clause built from \(lits\) together with this distinctness witness, then the list of variable indices of \(lits\) has no repeated entries.
Let \(lits\) be a list of Boolean-circuit literals over \(n\) variables, and suppose the list of variable indices \(\mathrm{idx}(l)\) taken over the literals \(l\) in \(lits\) has no repeated entry. Then the term obtained from the normal-form AND-clause on \(lits\) by mapping each circuit literal to its switching-lemma literal, viewed as a list of literals, has no repeated entry.
Let \(lits\) be a list of circuit literals on \(n\) variables whose variable indices are pairwise distinct, and let \(C\) be the AND-clause built from \(lits\). If \(\ell _1\) and \(\ell _2\) are any two literals occurring in the term obtained by converting \(C\), and \(\ell _1\) and \(\ell _2\) have the same variable, then \(\ell _1 = \ell _2\).
Fix \(n\), and let \(lits\) be a list of circuit literals on \(n\) variables whose variable indices are pairwise distinct. Form the normal-form AND-clause on this list, and convert it into a term by replacing each circuit literal with the corresponding switching-lemma literal. Then the width of that term — its number of literals — equals the length of \(lits\).
Let \(cs\) be a list of normal-form OR circuits on \(n\) variables, and suppose that every entry of \(cs\) is an OR-clause — that is, a base clause given by a list of literals, rather than an internal OR node. Then for every assignment \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\), evaluating the normal-form AND circuit that conjoins the circuits in \(cs\) agrees with evaluating the CNF formula built from it: the value at \(x\) of the AND-node over \(cs\) equals the value at \(x\) of the CNF whose clauses are obtained by converting each OR-clause of \(cs\) into a disjunction of literals.
Let \(cs\) be a list of normal-form OR circuits on \(n\) variables, and suppose every member of \(cs\) is an OR-clause \(\mathrm{clause}\, (lits)\) whose literals have pairwise distinct variable indices. Form the depth-2 AND circuit that conjoins the circuits in \(cs\), and pass to the CNF formula obtained by turning each of these OR-clauses into a clause (a disjunction of literals). Then every clause of this CNF is a list of literals with no repeated entry.
Let \(cs\) be a list of normal-form OR circuits on \(n\) variables, and suppose every member of \(cs\) is a base OR-clause: for each \(c \in cs\) there is a list of circuit literals \(lits\), whose variable indices are pairwise distinct, with \(c = \mathrm{clause}\, lits\). Form the CNF formula of the AND-node over \(cs\) by converting each of these OR-clauses into a term. Then within any single clause of this CNF, any two literals that share the same variable are identical — in particular they carry the same negation flag.
Fix \(n \in \mathbb {N}\), let \(cs\) be a list of normal-form OR circuits on \(n\) variables, and let \(w \in \mathbb {N}\). Suppose every circuit in \(cs\) is an OR-clause on some list of literals whose variable indices are pairwise distinct and whose length is at most \(w\). Form the normal-form AND circuit that conjoins all the members of \(cs\), and convert that circuit to a CNF formula by turning each OR-clause into a clause (a disjunction of its literals). Then this CNF formula has width at most \(w\); equivalently, each of its clauses contains at most \(w\) literals.
Recursively expands a node \(v\) of layer \(m\) of a Boolean feedforward circuit into a tree-shaped \(\texttt{Circuit}\ n\): a layer-\(0\) node becomes the positive literal of the corresponding input variable, and a node at a higher layer becomes an AND/OR node whose children are the unrollings of the nodes on its input wires. Nodes feeding several downstream gates are duplicated.
Let \(F\) be a Boolean feedforward circuit whose inputs are indexed by \(\mathrm{Fin}\, n\), and suppose each of its gates has finite arity and computes either the conjunction or the disjunction of its inputs, according to a flag marking every non-input node as an AND node or an OR node. Then for every layer \(m\) with \(m \le \mathrm{depth}(F)\), every node \(v\) of layer \(m\), and every input assignment \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\), the Boolean circuit obtained by tree-unrolling \(v\), evaluated at \(x\), equals the value that \(v\) takes in \(F\) under the assignment \(x\).
Let \(F\) be a Boolean feedforward circuit from the inputs \(\mathrm{Fin}\, n\) to an output type, in which every gate has finitely many input wires, and suppose \(k \in \mathbb {N}\) bounds these arities, so that each gate has at most \(k\) input wires. Then for every layer \(m \le \mathrm{depth}(F)\) and every node \(v\) of layer \(m\), the tree-unrolling of \(v\) is a Boolean circuit of size at most \((k+1)^m\).
Fix a natural number \(n\), and let \(\ell \) be a list of literals over \(n\) variables, each literal being a pair consisting of a variable index in \(\mathrm{Fin}\, n\) and a sign in \(\mathrm{Bool}\). Suppose the variable indices occurring in \(\ell \) are pairwise distinct, and let \(c\) be the OR-clause circuit built from \(\ell \) under this distinctness hypothesis. Then the list of variable indices of \(\ell \) is duplicate-free, i.e. its entries are pairwise distinct.
Let \(lits\) be a list of circuit literals in \(n\) variables whose underlying variable indices are pairwise distinct. Then the term obtained by converting the OR-clause on \(lits\) — the list of switching-lemma literals produced by sending each circuit literal to its counterpart — has no repeated entries.
Let \(n \ge 0\) and let \(lits\) be a list of circuit literals over \(n\) variables whose variable indices are pairwise distinct, and form the term obtained by converting each literal of \(lits\) to a switching-lemma literal (recording the same variable and polarity). Then any two literals of this term that share the same variable \(\mathtt{var} \in \mathrm{Fin}\, n\) are in fact equal.
Let \(\mathrm{lits}\) be a list of circuit literals over \(n\) variables whose variable indices are pairwise distinct. Then the width of the term obtained from the OR-clause on \(\mathrm{lits}\) — that is, the term formed by converting each circuit literal to its corresponding literal — equals the number of literals in \(\mathrm{lits}\).
Let \(cs\) be a list of normal-form AND circuits on \(n\) variables, and suppose each member of \(cs\) is a base AND-clause (a conjunction of literals with distinct variable indices) rather than an AND-node built from OR-subcircuits. Form the normal-form OR circuit whose children are exactly the circuits in \(cs\), so that it computes their disjunction. Then for every assignment \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\), this OR circuit evaluates at \(x\) to the same Boolean value as the DNF formula obtained from it by converting each child AND-clause into a term.
Let \(cs\) be a list of normal-form AND circuits on \(n\) variables, and suppose every element of \(cs\) is a clause — that is, each is given by a list of circuit literals whose variable indices are pairwise distinct. Consider the normal-form OR circuit that combines the members of \(cs\), and let its associated DNF formula be the list of terms obtained by converting each clause into a term (mapping every circuit literal to the corresponding switching-lemma literal with the same variable and polarity). Then every term of this DNF is duplicate-free: no term contains the same literal twice.
Fix \(n \in \mathbb {N}\) and a list \(c_1, \dots , c_m\) of normal-form AND circuits on \(n\) variables, and suppose each \(c_i\) is a base AND-clause — that is, \(c_i = \mathrm{clause}(\ell _1, \dots , \ell _{k_i})\) for some list of circuit literals whose variable indices are pairwise distinct. Let \(F\) be the DNF formula obtained from the OR circuit \(\mathrm{node}(c_1, \dots , c_m)\) by converting each clause \(c_i\) into the term formed by its literals. Then in every term of \(F\), any two literals that share the same variable are identical.
Let \(cs\) be a finite list of normal-form AND circuits on \(n\) variables, and let \(w \in \mathbb {N}\). Suppose every circuit in \(cs\) is a base clause — a list of literals whose variable indices are pairwise distinct — and that each such clause contains at most \(w\) literals. Then the DNF obtained from the normal-form OR circuit whose disjuncts are the circuits in \(cs\), by converting each AND-clause into a term, has width at most \(w\); that is, every term in it has at most \(w\) literals.
Let \(c\) be a Boolean circuit on \(n\) variables — a tree of literals combined by AND and OR gates — and let \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\) be an assignment to those variables. Converting \(c\) into strictly alternating AND-normal form and evaluating the resulting normal-form AND circuit at \(x\) yields the same Boolean value as evaluating \(c\) directly at \(x\).
Let \(c\) be a Boolean circuit on \(n\) variables, and let \(c'\) be the normal-form AND circuit obtained by normalizing it (a literal becomes a singleton clause, an AND gate normalizes each child to OR-normal form, and an OR gate is wrapped as a single OR node over the AND-normalizations of its children). Then \(c'\) contains exactly as many literal occurrences as \(c\): the number of literals in \(c'\) equals the number of literals in \(c\).
Let \(c\) be a Boolean circuit on \(n\) variables, where the size of a circuit is its total number of nodes, counting internal AND/OR gates together with literal leaves. Normalizing \(c\) into strictly alternating AND-normal form — a literal becoming a singleton clause, an AND gate normalizing each child into OR-normal form, and an OR gate becoming a single wrapping OR node over its children’s AND-normal forms — yields a normal-form AND circuit whose size is at most twice the size of \(c\).
Let \(c\) be a Boolean circuit on \(n\) variables and let \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\) be an assignment. Then both alternating normal forms of \(c\) compute the same value as \(c\) itself: the normal-form AND circuit obtained by normalizing \(c\) evaluates at \(x\) to \(c\)’s value, and likewise the normal-form OR circuit obtained by normalizing \(c\) evaluates at \(x\) to \(c\)’s value.
Let \(c\) be a Boolean circuit on \(n\) variables, whose literal count is the number of literal occurrences it contains (a leaf literal counting once, and an AND or OR gate counting the total over its children). Converting \(c\) to strictly alternating AND/OR normal form preserves this count: both the AND-normal form of \(c\) and the OR-normal form of \(c\) have the same number of literal occurrences as \(c\) itself.
Let \(c\) be a Boolean circuit on \(n\) variables, and write \(\mathrm{size}(c)\) for its total node count. Converting \(c\) into either alternating normal form increases its size by at most a factor of two: the normal-form AND circuit obtained by normalizing \(c\) has size at most \(2\, \mathrm{size}(c)\), and the normal-form OR circuit obtained by normalizing \(c\) likewise has size at most \(2\, \mathrm{size}(c)\).
Let \(c\) be a Boolean circuit on \(n\) variables and let \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\) be an assignment. Then the normal-form OR circuit obtained by OR-normalizing \(c\) computes the same Boolean value as \(c\) itself: evaluated at \(x\), it agrees with the value of \(c\) at \(x\).
For every Boolean circuit \(c\) on \(n\) variables, converting \(c\) to OR-normal form leaves its total number of literal occurrences unchanged: the literal count of the normal-form OR circuit obtained from \(c\) equals the literal count of \(c\) itself.
Let \(c\) be a Boolean circuit on \(n\) variables, and let \(\mathrm{size}(c)\) denote its total node count (literal leaves plus internal AND/OR gates). Write \(c'\) for the OR-normal form of \(c\), the normal-form OR circuit obtained by rewriting \(c\) into strictly alternating AND/OR normal form. Then the size of \(c'\) is at most twice the size of \(c\):
A Boolean circuit on \(n\) variables is a tree whose leaves are literals (lit \(l\)) and whose internal nodes (node \(isAnd\) \(children\)) are AND gates when \(isAnd = \mathrm{true}\) and OR gates when \(isAnd = \mathrm{false}\), applied to a list of child circuits. No alternation or deduplication constraint is imposed.
Let \(P\) be a property of Boolean circuits on \(n\) variables. Suppose that \(P\) holds for every literal circuit, and that whenever \(P\) holds for every circuit in a list \(cs\) of children, it also holds for the gate circuit \(\mathrm{node}(isAnd, cs)\) formed from \(cs\) (an AND gate when \(isAnd = \mathrm{true}\), an OR gate when \(isAnd = \mathrm{false}\)). Then \(P\) holds for every Boolean circuit on \(n\) variables.
Given a circuit \(c\) on \(m\) variables and a function \(f : \mathrm{Fin}\, m \to \mathrm{Fin}\, m'\), the re-indexed circuit \(\mathrm{reidx}(c, f)\) on \(m'\) variables is obtained by replacing each literal index \(i\) by \(f(i)\) (keeping its sign) and recursing structurally through the \(\mathrm{And}/\mathrm{Or}\) nodes.
Let \(c\) be a Boolean circuit on \(m\) variables and let \(f : \mathrm{Fin}\, m \to \mathrm{Fin}\, m'\) be any map of variable indices. Then re-indexing \(c\) along \(f\) leaves its depth unchanged: \(\mathrm{depth}(\mathrm{reidx}(c, f)) = \mathrm{depth}(c)\).
Let \(c\) be a Boolean circuit on \(m\) variables, let \(f : \mathrm{Fin}\, m \to \mathrm{Fin}\, m'\) be a re-indexing map, and let \(g : \mathrm{Fin}\, m' \to \mathbb {F}_2\) be an assignment to the \(m'\) variables. Then evaluating the re-indexed circuit \(\mathrm{reidx}(c, f)\) under \(g\) gives the same value as evaluating the original circuit \(c\) under the pulled-back assignment \(g \circ f\):
Turns a circuit \(C\) on \(n\) inputs into a feedforward circuit \(\texttt{FeedForward}\ \texttt{Bool}\ (\mathrm{Fin}\, n)\ \mathrm{Unit}\) of depth \(C.\texttt{depth} + 1\): layer \(0\) carries the \(n\) input variables and every higher layer carries a single wire. The gate from layer \(0\) to layer \(1\) computes \(C.\texttt{eval}\) from all inputs at once, and all later gates are identity wires that pass the single Boolean value upward.
Let \(C\) be a Boolean circuit on \(n\) variables, with depth \(\mathrm{depth}(C)\) equal to the length of the longest root-to-leaf path in \(C\). Embedding \(C\) as a layered feedforward circuit — in which layer \(0\) carries the \(n\) input variables and each higher layer carries a single wire — produces a circuit whose depth equals \(\mathrm{depth}(C) + 1\).
Let \(C\) be a Boolean circuit on \(n\) variables, and let \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\) be an assignment of Boolean values to those variables. Form the feedforward embedding of \(C\), a layered feedforward circuit with a single output wire, and let its unique output value under \(x\) be the Boolean it computes. Then this value equals \(C(x)\), the value obtained by evaluating the circuit \(C\) directly on the assignment \(x\).
Let \(C\) be a Boolean circuit on \(n\) variables and let \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\) be an assignment of Boolean values to the inputs of \(C\). The feedforward embedding of \(C\) is a layered circuit of depth \(\mathrm{depth}(C) + 1\) whose layer \(0\) carries the \(n\) input variables and each of whose higher layers carries a single node. Then for every layer index \(m\) with \(1 \le m \le \mathrm{depth}(C) + 1\) and every node \(v\) of layer \(m\), the value assigned to \(v\) by the feedforward embedding on input \(x\) equals \(\mathrm{eval}(C, x)\), the value of \(C\) on \(x\).
Let \(C\) be a Boolean circuit on \(n\) variables, and let \(\mathrm{size}(C)\) and \(\mathrm{depth}(C)\) denote its total number of nodes and its depth, respectively. Writing \(\widehat{C}\) for the layered feedforward circuit obtained by embedding \(C\), the number of non-input nodes of \(\widehat{C}\) satisfies
Converts a general circuit into a normal-form AND circuit: a literal becomes a singleton clause; an AND node maps its children through \(\mathrm{toNOr}\); and an OR node is wrapped as a single OR node built from the children’s \(\mathrm{toNAnd}\) images, enforcing alternation.
Converts a general circuit into a normal-form OR circuit: a literal becomes a singleton clause; an OR node maps its children through \(\mathrm{toNAnd}\); and an AND node is wrapped as a single AND node built from the children’s \(\mathrm{toNOr}\) images, enforcing alternation.
Under the sign encoding \(\sigma \colon \mathrm{Bool}\to \{ -1,1\} \), the Boolean value \(\mathrm{false}\) is sent to \(1\); that is, \(\sigma (\mathrm{false}) = 1\).
Under the sign encoding \(\sigma \colon \mathrm{Bool}\to \{ -1,1\} \), which sends the truth value false to \(1\) and true to \(-1\), the truth value true is mapped to \(-1\); that is, \(\sigma (\text{true}) = -1\).
For every Boolean value \(b\), the sign encoding satisfies \(\sigma (b)^2 = 1\).
For every \(b \in \mathrm{Bool}\), the sign encoding \(\sigma (b) \in \{ -1,1\} \) satisfies \(\sigma (b)\cdot \sigma (b) = 1\); that is, \(\sigma (b)^2 = 1\).
Let \(\sigma \colon \mathrm{Bool}\to \{ -1,1\} \) be the sign encoding, so that \(\sigma (\mathrm{false})=1\) and \(\sigma (\mathrm{true})=-1\). Then for every Boolean value \(b\), negating \(b\) flips the sign of its encoding:
The Walsh–Fourier character associated to \(S\subseteq [n]\) is
The family \(\{ \chi _S\} _{S\subseteq [n]}\) forms an orthonormal basis for \(L^2(\{ 0,1\} ^n,\mathrm{uniform})\).
Fix a subset \(S\subseteq [n]\) of coordinates. For every point \(x\) of the Boolean hypercube \(\{ 0,1\} ^n\) and every coordinate \(i\in [n]\), the value of the Walsh–Fourier character \(\chi _S\) at the point \(x^i\) obtained by flipping the \(i\)-th coordinate of \(x\) is
Fix \(n\), and for a subset \(S\subseteq [n]\) let \(\chi _S\colon \{ 0,1\} ^n\to \mathbb {R}\) denote the associated Walsh–Fourier character. Then for all subsets \(S,T\subseteq [n]\) and every point \(x\in \{ 0,1\} ^n\),
where \(S\, \triangle \, T\) is the symmetric difference of \(S\) and \(T\).
Let \(n\) be a natural number, and let \(S, T \subseteq [n]\) be disjoint subsets of coordinates. Then for every point \(x\) of the Boolean hypercube \(\{ 0,1\} ^n\), the Walsh–Fourier character of the union factors as the product of the characters of the two sets:
Let \(n\) be a natural number, let \(S \subseteq [n]\), and let \(x \in \{ 0,1\} ^n\). Writing \(\lnot x\) for the point obtained by flipping every coordinate of \(x\), the Walsh–Fourier character \(\chi _S\) satisfies
where \(\left\lvert S\right\rvert \) is the number of elements of \(S\).
Fix a dimension \(n\), a coordinate \(i \in [n]\), and a point \(x\) of the Boolean hypercube \(\{ 0,1\} ^n\). Then the Walsh–Fourier character of the singleton set \(\{ i\} \), evaluated at \(x\), is the sign encoding of the \(i\)-th coordinate of \(x\):
Let \(n\) be a natural number, let \(S \subseteq [n]\), and let \(x \in \{ 0,1\} ^n\) be a point of the Boolean hypercube. Then the Walsh–Fourier character \(\chi _S\) satisfies \(\chi _S(x)^2 = 1\); that is, every character takes values in \(\{ -1, 1\} \).
Fix an arity \(n\) and a coordinate direction \(i \in \{ 1, \dots , n\} \). For any two Boolean functions \(f, g : \{ 0,1\} ^n \to \mathbb {R}\), the discrete derivative in direction \(i\) satisfies
that is, the operator \(D_i\) is additive.
Fix an arity \(n\) and a coordinate direction \(i \in \{ 1,\dots ,n\} \). For every scalar \(c \in \mathbb {R}\) and every Boolean function \(f : \{ 0,1\} ^n \to \mathbb {R}\), the discrete derivative in direction \(i\) commutes with scalar multiplication:
that is, scaling \(f\) by \(c\) scales its discrete derivative by the same factor.
For each coordinate \(i \in [n]\), the dictator function on the Boolean hypercube \(\{ 0,1\} ^n\) coincides with the Walsh–Fourier character associated to the singleton set \(\{ i\} \); that is, \(\mathrm{dict}_i = \chi _{\{ i\} }\) as functions \(\{ 0,1\} ^n \to \mathbb {R}\).
Fix \(n \in \mathbb {N}\) and a coordinate direction \(i \in \{ 1,\dots ,n\} \), and write \(E_i\) for the expectation operator in direction \(i\), which sends a Boolean function \(h : \{ 0,1\} ^n \to \mathbb {R}\) to the function \((E_i h)(x) = \tfrac {1}{2}\bigl(h(x^{i\mapsto 0}) + h(x^{i\mapsto 1})\bigr)\). Then \(E_i\) is additive: for all Boolean functions \(f, g : \{ 0,1\} ^n \to \mathbb {R}\),
Let \(n\) be an arity and let \(i\) be one of its coordinates. For every scalar \(c \in \mathbb {R}\) and every Boolean function \(f : \{ 0,1\} ^n \to \mathbb {R}\), applying the expectation operator in direction \(i\) to the scaled function \(c\, f\) yields \(c\) times the expectation operator applied to \(f\):
Let \(x\in \{ 0,1\} ^n\) be a point of the Boolean hypercube, and let \(i,j\) be coordinate indices with \(i\ne j\). Writing \(x^i\) for the point obtained from \(x\) by flipping the \(i\)-th coordinate, we have \((x^i)_j = x_j\); that is, flipping coordinate \(i\) leaves coordinate \(j\) unchanged.
For any two subsets \(S, T \subseteq [n]\), the \(L^2\) inner product of the Walsh–Fourier characters \(\chi _S\) and \(\chi _T\), taken with respect to the uniform measure on \(\{ 0,1\} ^n\), satisfies
In particular, the family \(\{ \chi _S\} _{S\subseteq [n]}\) is orthonormal.
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be an odd Boolean function, meaning \(f(\lnot x) = -f(x)\) for every \(x \in \{ 0,1\} ^n\), where \(\lnot x\) denotes the pointwise negation of \(x\). Then for every subset \(S \subseteq [n]\) of even cardinality, the Fourier–Walsh coefficient of \(f\) at \(S\) vanishes: \(\hat f(S) = 0\).
Let \(X\) be measurable on a probability space, assume \(X\in L^4\), \(X^2\) and \(X^4\) are integrable, and \(\mathbb {E}X^2\gt 0\). If \(B\ge 1\) and \(\mathbb {E}X^4\le B(\mathbb {E}X^2)^2\), then for every \(0\le t\le 1\), \(\Pr [X^2\gt t^2\mathbb {E}X^2]\ge (1-t^2)^2/B\).
Let \(X\in L^4\) on a probability space, let \(B\ge 1\), and suppose \(\mathbb {E}X^4\le B(\mathbb {E}X^2)^2\), \(X^4\) is integrable, and \(\mathbb {E}X^2\gt 0\). For every \(t\gt 0\), \(\Pr [X^4\ge t^4(\mathbb {E}X^2)^2]\le B/t^4\).
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function of degree at most \(k\), meaning that every nonzero Fourier coefficient \(\hat f(S)\) is supported on a set \(S\) with \(|S| \le k\). Writing \(\mathbb {E}\) for the expectation under the uniform measure on \(\{ 0,1\} ^n\), we then have
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function whose Fourier expansion has degree at most \(k\); that is, every nonzero Fourier coefficient \(\hat f(S)\) is supported on a set \(S\) with \(\left\lvert S\right\rvert \le k\). When \(\{ 0,1\} ^n\) is equipped with the uniform probability measure, \(f\) is \(9^k\)-reasonable: its fourth moment is controlled by the square of its second moment via
On the one-dimensional Boolean hypercube \(\{ 0,1\} ^1\), whose points are the functions from a single-element index set to \(\{ 0,1\} \), the set of all points consists of exactly the two constant functions: the one taking the value \(0\) (false) and the one taking the value \(1\) (true).
Fix a natural number \(n\). Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function, and let \(p,u\in \mathbb {R}\) satisfy \(1\le p\le 2\le u\). Set \(\rho =\sqrt{(p-1)/(u-1)}\) and let \(T_\rho \) denote the noise operator with parameter \(\rho \). Then
where the expectations are taken with respect to the uniform measure on \(\{ 0,1\} ^n\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(S \subseteq \{ 0, 1, \dots , n-1\} \) be a finite set of indices, and let \(\iota \colon \{ 0, 1, \dots , n-1\} \hookrightarrow \{ 0, 1, \dots , n\} \) be the order-preserving inclusion that regards each index as an element of the enlarged range. Then the image \(\iota (S)\) has the same cardinality as \(S\), that is, \(\left\lvert \iota (S)\right\rvert = \left\lvert S\right\rvert \).
Let \(n\) be a natural number and let \(S \subseteq \{ 0, 1, \dots , n-1\} \). Embed \(S\) into \(\{ 0, 1, \dots , n\} \) by the map that sends each index to itself, and adjoin the top index \(n\). Since \(n\) is not the image of any element of \(S\), this union satisfies
Let \(n\ge 0\) be an integer, let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be functions, and write \(\mathbb {E}[h] = 2^{-n}\sum _{x \in \{ 0,1\} ^n} h(x)\) for the expectation under the uniform measure and \(\langle f, g\rangle = \mathbb {E}[f\cdot g]\) for the \(L^2\) inner product. Then
Fix \(n\in \mathbb {N}\), and let \(S\subseteq [n]\), \(x\in \{ 0,1\} ^n\), and \(b\in \{ 0,1\} \). Form the point \((x,b)\in \{ 0,1\} ^{n+1}\) obtained by appending \(b\) as the final coordinate, and regard \(S\) as a subset \(S'\subseteq [n+1]\) via the inclusion \([n]\hookrightarrow [n+1]\) that leaves each index fixed (so the new last coordinate never lies in \(S'\)). Then the Walsh–Fourier character of \(S'\) evaluated at the extended point agrees with that of \(S\) at the original point:
Let \(S \subseteq [n]\), let \(x \in \{ 0,1\} ^n\), and let \(b\) be a bit. Write \((x,b) \in \{ 0,1\} ^{n+1}\) for the point obtained by appending \(b\) as the final coordinate, and let \(S' \subseteq [n+1]\) be the image of \(S\) under the inclusion \([n] \hookrightarrow [n+1]\) of the first \(n\) coordinates, together with the new last coordinate \(n+1\). Then
where \(\sigma (0)=1\) and \(\sigma (1)=-1\); that is, adjoining the last coordinate to the lifted index set multiplies the Walsh character by the sign of \(b\).
Fix \(n\), let \(f\colon \{ 0,1\} ^n\to \mathbb {R}\) be any Boolean function, and let \(\rho \in \mathbb {R}\) satisfy \(\rho ^2\le 1\). Writing \(T_\rho \) for the noise operator, so that \(T_\rho f(x)=\sum _{S\subseteq [n]}\rho ^{|S|}\, \hat f(S)\, \chi _S(x)\), one has
where the expectation is taken with respect to the uniform measure on \(\{ 0,1\} ^n\). Here the Walsh characters are \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and the Fourier coefficients are \(\widehat f(S)=\mathbb {E}[f\chi _S]\), with uniform expectation.
Fix a dimension \(n\) and a correlation parameter \(\rho \) with \(0 \le \rho \le 1\), and let \(K_\rho (x,y) = \prod _{i} \tfrac {1 + \rho \, \chi (x_i)\, \chi (y_i)}{2}\) denote the noise kernel on the Boolean hypercube \(\{ 0,1\} ^n\), where \(\chi (0)=1\) and \(\chi (1)=-1\). If \(h, h' : \{ 0,1\} ^n \times \{ 0,1\} ^n \to \mathbb {R}\) satisfy \(h(x,y) \le h'(x,y)\) for all \(x\) and \(y\), then the kernel-weighted average of \(h\) is at most that of \(h'\):
Let \(f\colon \{ 0,1\} ^{n+1}\to \mathbb {R}\) be a Boolean function on \(n+1\) variables, and suppose \(f\) has degree at most \(k\); that is, every set \(S\) with nonzero Fourier coefficient \(\hat f(S)\ne 0\) satisfies \(\left\lvert S\right\rvert \le k\). Form the average of \(f\) over its last coordinate, the function on \(n\) variables obtained by fixing that coordinate to each of its two values and averaging,
Then this averaged function also has degree at most \(k\): each set \(S\) with a nonzero Fourier coefficient of the average satisfies \(\left\lvert S\right\rvert \le k\).
Let \(f\colon \{ 0,1\} ^{n+1}\to \mathbb {R}\) be a Boolean function of degree at most \(k\), meaning that every frequency set \(S\) with nonzero Fourier–Walsh coefficient \(\hat f(S)\ne 0\) satisfies \(\left\lvert S\right\rvert \le k\). Write \(f_0\) and \(f_1\) for the two restrictions of \(f\) to \(\{ 0,1\} ^{n}\) obtained by fixing its last coordinate to \(0\) and to \(1\) respectively, and let \(g=\tfrac 12(f_0-f_1)\) be the half-difference of \(f\) over that coordinate. Then \(g\) is a Boolean function on \(\{ 0,1\} ^{n}\) of degree at most \(k-1\), with \(k-1\) interpreted as \(0\) when \(k=0\).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function of degree at most \(0\); that is, every set \(S\) with nonzero Fourier–Walsh coefficient \(\hat f(S)\ne 0\) satisfies \(|S|\le 0\). Then \(f\) is constant: for every point \(x\in \{ 0,1\} ^n\) one has \(f(x)=f(x_0)\), where \(x_0\) is the all-\(\mathtt{false}\) point.
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function of degree at most \(0\), meaning that every set \(S\) with a nonzero Fourier coefficient \(\hat f(S)\ne 0\) satisfies \(|S|\le 0\). Then, with expectations taken under the uniform measure on \(\{ 0,1\} ^n\),
Let \(f\colon \{ 0,1\} ^1\to \mathbb {R}\) be a Boolean function of one bit and let \(p\in \mathbb {R}\). Writing \(\hat f(\emptyset )\) and \(\hat f(\{ 0\} )\) for its two Fourier–Walsh coefficients, the expectation of \(|f|^p\) under the uniform measure is
Here real powers use \(0^p=0\) for \(p\ne 0\) and \(0^0=1\).
Let \(g, h : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions, and let \(\mathbb {E}[\, \cdot \, ]\) denote the expectation under the uniform measure on the hypercube \(\{ 0,1\} ^n\). Then
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function on the Boolean hypercube. Then its fourth moment under the uniform measure is nonnegative:
Let \(f\colon \{ 0,1\} ^1\to \mathbb {R}\) be a Boolean function of arity one, and let \(\rho \) be any real number. Define \(T_\rho \) by multiplying the Fourier coefficient at each subset \(S\) of the coordinate set by \(\rho ^{|S|}\); then its expected square under the uniform measure is
where \(\emptyset \) and \(\{ 0\} \) are the only two frequencies for a single coordinate. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function and let \(p\in \mathbb {R}\) be any real exponent. Using \(0^0=1\) and \(0^p=0\) for \(p\ne 0\), including negative \(p\), the expected value of \(\left\lvert f\right\rvert ^p\) under the uniform measure on the hypercube is nonnegative:
Let \(f\colon \{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function and let \(\rho \in \mathbb {R}\) be arbitrary. Applying the noise operator \(T_\rho \) to \(f\) and averaging the square over the uniform measure gives a nonnegative quantity:
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
For any two Boolean functions \(g, h : \{ 0,1\} ^n \to \mathbb {R}\), the expectation of the product of their squares under the uniform measure is nonnegative:
Let \(\varphi : \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a Boolean function on \(n+1\) variables, and for a bit \(b \in \{ 0,1\} \) let \(\varphi _b : \{ 0,1\} ^n \to \mathbb {R}\), \(\varphi _b(x) = \varphi (x, b)\), be its restriction obtained by fixing the last coordinate to \(b\). Then the expectation of \(\varphi \) under the uniform measure on \(\{ 0,1\} ^{n+1}\) is the average of the expectations of its two restrictions:
Let \(n\) be a natural number and let \(h : \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a Boolean function on \(n+1\) bits. Writing \((x', b) \in \{ 0,1\} ^n \times \{ 0,1\} \) for the point of \(\{ 0,1\} ^{n+1}\) obtained by appending the bit \(b\) to \(x'\) as its last coordinate, the uniform expectation of \(h\) can be computed by first averaging over the last bit and then over the remaining \(n\) bits:
Fix \(n \in \mathbb {N}\) and identify a subset of \(\{ 0, 1, \dots , n\} \) that omits the top element \(n\) with the corresponding subset of \(\{ 0, 1, \dots , n-1\} \), via the inclusion \(\iota \) that sends \(T \subseteq \{ 0,\dots ,n-1\} \) to the same set regarded inside \(\{ 0,\dots ,n\} \). Then for every function \(\varphi \) assigning a real number to each subset of \(\{ 0, 1, \dots , n\} \),
the two terms collecting the subsets of \(\{ 0,\dots ,n\} \) that exclude and that include the element \(n\), respectively.
Let \(f\colon \{ 0,1\} ^{n+1}\to \mathbb {R}\), let \(g\) be the average of \(f\) over its last coordinate — the Boolean function of arity \(n\) given by \(g(x) = \tfrac 12\bigl(f(x,\text{false}) + f(x,\text{true})\bigr)\) — and let \(S \subseteq [n]\) be a set of coordinates among the first \(n\). Write \(\iota \colon [n] \hookrightarrow [n+1]\) for the inclusion sending each of the first \(n\) coordinates to itself and omitting the last, so that \(\iota (S)\) is \(S\) viewed as a subset of \([n+1]\) that avoids the coordinate \(n+1\). Then the Fourier–Walsh coefficient of \(g\) at frequency \(S\) equals the Fourier–Walsh coefficient of \(f\) at frequency \(\iota (S)\):
Let \(f\colon \{ 0,1\} ^{n+1}\to \mathbb {R}\) be a Boolean function on \(n+1\) variables, and let \(g\colon \{ 0,1\} ^{n}\to \mathbb {R}\) be its half-difference over the last coordinate, given by \(g(x)=\tfrac 12\bigl(f(x,0)-f(x,1)\bigr)\). For every frequency set \(S\subseteq \{ 1,\dots ,n\} \), the Fourier–Walsh coefficient of \(g\) at \(S\) equals the Fourier–Walsh coefficient of \(f\) at the frequency set obtained by regarding \(S\) as a subset of the first \(n\) coordinates of \(\{ 1,\dots ,n+1\} \) and adjoining the last coordinate:
Let \(f : \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a Boolean function on \(n+1\) variables, and let \(g\) and \(h\) be the average and the half-difference of \(f\) over its last coordinate, so that \(g(x) = \tfrac 12\bigl(f(x,0) + f(x,1)\bigr)\) and \(h(x) = \tfrac 12\bigl(f(x,0) - f(x,1)\bigr)\) for \(x \in \{ 0,1\} ^n\). Then, with all expectations taken under the uniform measure, the fourth moment of \(f\) decomposes as
Let \(f:\{ 0,1\} ^{n+1}\to \mathbb {R}\) be a Boolean function, and let \(\rho \in \mathbb {R}\). Write \(g\) for the average of \(f\) over its last coordinate and \(h\) for the corresponding half-difference, so that
are Boolean functions of the first \(n\) variables. Then the fourth moment of \(T_\rho f\) under the uniform measure decomposes as
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix a natural number \(n\). Let \(1 \le p \le u\) with \(u \gt 1\), and let \(\rho \) satisfy \(0 \le \rho \le 1\) together with \(\rho \le \sqrt{(p-1)/(u-1)}\). Then for every Boolean function \(f:\{ 0,1\} ^n \to \mathbb {R}\), the noise operator \(T_\rho \) contracts the \(L^u\) norm below the \(L^p\) norm:
where the expectations are taken over the uniform measure on \(\{ 0,1\} ^n\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix a natural number \(n\), real exponents \(1 \le p \le u\) with \(u \ge 2\), and a noise rate \(\rho \) with \(0 \le \rho \le 1\) satisfying \(\rho \le \sqrt{(p-1)/(u-1)}\). Then for every pair of Boolean functions \(f, g : \{ 0,1\} ^n \to \mathbb {R}\),
where \(\langle \cdot ,\cdot \rangle \) is the \(L^2\) inner product, \(\mathbb {E}\) the expectation, and \(T_\rho \) the noise operator, all taken with respect to the uniform measure on the hypercube. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix a natural number \(n\). Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function, and let \(p, u\) be real exponents with \(2 \le p \le u\). Suppose \(\rho \in \mathbb {R}\) satisfies \(0 \le \rho \le 1\) together with \(\rho \le \sqrt{(p-1)/(u-1)}\). Then, with expectations taken under the uniform measure on \(\{ 0,1\} ^n\) and \(T_\rho \) the noise operator, the higher \(L^u\) norm of the noised function is bounded by the lower \(L^p\) norm of \(f\):
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(n\) be a natural number, let \(p \gt 1\) and let \(f, h : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions, with expectations taken under the uniform measure on \(\{ 0,1\} ^n\). Then their \(L^2\) inner product satisfies
where the exponents \(p\) and \(p/(p-1)\) are conjugate.
Fix an integer dimension \(n\ge 0\) and real Hölder conjugate exponents \(p, q\) (so \(p, q \gt 1\) and \(\tfrac 1p + \tfrac 1q = 1\)). For a function \(g : \{ 0,1\} ^n \to \mathbb {R}\) write \(\lVert g\rVert _r = \bigl(\mathbb {E}\bigl[\, |g|^r\, \bigr]\bigr)^{1/r}\) for its \(L^r\) norm under the uniform measure on \(\{ 0,1\} ^n\). Then for every function \(u : \{ 0,1\} ^n \to \mathbb {R}\) there exists a function \(f : \{ 0,1\} ^n \to \mathbb {R}\) with \(\lVert f\rVert _p \le 1\) and \(\lVert u\rVert _q \le \langle f, u\rangle \), where \(\langle f, u\rangle = \mathbb {E}[f\cdot u]\).
Let \(\rho \in \mathbb {R}\) satisfy \(\rho ^2\le 1\), and let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function. Then the noise operator \(T_\rho \) does not increase the \(L^2\) norm under the uniform measure:
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix an integer \(k \ge 1\) and a real number \(\rho \) with \(\rho ^2 \le \dfrac {1}{2k-1}\). Then for every function \(f : \{ 0,1\} ^n \to \mathbb {R}\) on the Boolean hypercube,
where the expectation is taken under the uniform measure on \(\{ 0,1\} ^n\) and \(T_\rho \) is the noise operator \(T_\rho f = \sum _{S \subseteq [n]} \rho ^{|S|}\, \hat f(S)\, \chi _S\). Equivalently, \(\lVert T_\rho f\rVert _{2k} \le \lVert f\rVert _2\). Here the Walsh characters are \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and the Fourier coefficients are \(\widehat f(S)=\mathbb {E}[f\chi _S]\), with uniform expectation.
Fix an integer \(k \ge 1\) and a real noise rate \(\rho \) satisfying \(\rho ^2 \le \tfrac {1}{2k-1}\). Then for every Boolean function \(f\colon \{ 0,1\} ^n\to \mathbb {R}\), the image \(T_\rho f\) under the noise operator obeys the \(L^2\)-to-\(L^{2k}\) hypercontractive bound
where the expectation is taken with respect to the uniform measure on \(\{ 0,1\} ^n\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function and let \(\rho \in \mathbb {R}\) be a noise parameter satisfying \(\rho ^2\le 1/3\). Then the fourth moment of the noised function is controlled by the square of its second moment:
where \(T_\rho \) is the noise operator and \(\mathbb {E}\) denotes expectation under the uniform measure on \(\{ 0,1\} ^n\). Equivalently, \(\lVert T_\rho f\rVert _4\le \lVert f\rVert _2\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(\rho \in \mathbb {R}\) satisfy \(\rho ^2\le \tfrac 15\), and let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be any Boolean function. Then the sixth moment of the noised function \(T_\rho f\) is controlled by the cube of the second moment of \(f\):
where both expectations are taken with respect to the uniform measure on \(\{ 0,1\} ^n\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(q \ge 2\) be an even integer, and let \(\rho \in \mathbb {R}\) satisfy \(\rho ^2 \le \tfrac {1}{q-1}\). Then for every Boolean function \(f : \{ 0,1\} ^n \to \mathbb {R}\), the noise operator \(T_\rho \) contracts from \(L^2\) to \(L^q\) in the sense that
where both expectations are taken with respect to the uniform measure on \(\{ 0,1\} ^n\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
For every Boolean function \(f:\{ 0,1\} ^n\to \mathbb {R}\), applying the noise operator with parameter \(\rho =1/\sqrt3\) contracts the \(L^2\) norm below the \(L^{4/3}\) norm of the original function:
where both expectations are taken under the uniform measure on the hypercube. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix exponents \(p, q \ge 1\) and a noise parameter \(\rho \) with \(0 \le \rho \le 1\), and write \(T_\rho \) for the associated noise operator. Suppose the two-function bound
holds for every pair of Boolean functions \(f, g \colon \{ 0,1\} ^1 \to \mathbb {R}\) on the one-dimensional cube. Then, for every \(n\), the same inequality holds for every pair of Boolean functions \(f, g \colon \{ 0,1\} ^n \to \mathbb {R}\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(1 \lt p\) and \(q \ge 2\) be conjugate exponents, so that \(1/p + 1/q = 1\), let \(\rho \in \mathbb {R}\), and let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function; write \(\lVert g\rVert _r = (\mathbb {E}[\left\lvert g\right\rvert ^r])^{1/r}\) for the \(L^r\) norm under the uniform measure on the hypercube. Suppose Hölder’s inequality holds for the exponent pair \((p,q)\), that is, \(\langle u, v\rangle \le \lVert u\rVert _p\, \lVert v\rVert _q\) for all Boolean functions \(u,v\), and suppose the noise operator satisfies the \((2,q)\) bound \(\mathbb {E}\bigl[\left\lvert T_\rho g\right\rvert ^q\bigr] \le \bigl(\mathbb {E}[g^2]\bigr)^{q/2}\) for every Boolean function \(g\). Then \(T_\rho \) satisfies the dual \((p,2)\) bound
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
For a Boolean function \(f : \{ 0,1\} ^n \to \mathbb {R}\), the \(L^2\) inner product of \(f\) with itself equals the expectation of its square: \(\langle f, f\rangle = \mathbb {E}[f^2]\), that is, \(2^{-n}\sum _{x \in \{ 0,1\} ^n} f(x)^2\).
Let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions, and take all expectations with respect to the uniform measure on \(\{ 0,1\} ^n\). Then their \(L^2\) inner product satisfies
Fix a dimension \(n\) and a real number \(\rho \). For any two Boolean functions \(f, g : \{ 0,1\} ^n \to \mathbb {R}\), the \(L^2\) inner product of \(f\) with the noised function \(T_\rho g\) can be written as a kernel-weighted double sum:
where \(K_\rho (x,y) = \prod _{i} \tfrac {1 + \rho \, \sigma (x_i)\, \sigma (y_i)}{2}\) is the noise kernel and \(\sigma \) is the \(\pm 1\) sign encoding of a bit. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(p,q,b\in \mathbb {R}\) satisfy \(1\le p\le q\le 2\) and \(0\le b\le 1\). Set \(\rho =\sqrt{(p-1)/(q-1)}\) when \(q\gt 1\) and \(\rho =0\) when \(q=1\). Then
Let \(X\colon \Omega \to \mathbb {R}\) be a function on a measurable space, let \(P\) be a measure on \(\Omega \), and let \(B\in \mathbb {R}\). The function \(X\) is \(B\)-reasonable under \(P\) when \(P\) is a probability measure, \(B\ge 1\), \(X\) is almost everywhere measurable with finite fourth absolute moment, and
Fix a dimension \(n\) and real exponents \(p, u\) with \(1 \lt p \le u \le 2\), and set the noise rate
Then for every function \(f : \{ 0,1\} ^n \to \mathbb {R}\) on the Boolean hypercube, the noise operator \(T_\rho \) contracts \(L^p\) into \(L^u\):
where the expectations are taken with respect to the uniform measure on \(\{ 0,1\} ^n\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(p, u, \rho \) be real numbers with \(1 \lt p \lt u \lt 2\), and suppose \(\rho ^2 = (p-1)/(u-1)\). Then there exist real numbers \(\theta \in (0,1)\) and \(s \gt 0\) such that
Let \(1 \lt p \le q \le 2\), let \(f : \{ 0,1\} ^1 \to \mathbb {R}\) be a Boolean function of arity \(1\), and set \(\rho = \sqrt{(p-1)/(q-1)}\). Then, with expectations taken under the uniform measure and \(T_\rho \) the noise operator,
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(n\ge 0\) be an integer, let \(f\colon \{ 0,1\} ^n\to \mathbb {R}\) be a function, and let \(r,s\) be real numbers with \(1\le r\le s\). Writing \(\mathbb {E}\) for the expectation under the uniform probability measure on \(\{ 0,1\} ^n\), the \(L^r\) norm of \(f\) is bounded by its \(L^s\) norm:
Let \(\Omega \) be a finite sample space with all subsets measurable, let \(\pi \) be a probability mass function on \(\Omega \), and give \(\Omega \) the probability measure \(P(\{ \omega \} )=\pi (\omega )\). For any function \(X\colon \Omega \to \mathbb {R}\), if \(\mu \gt 0\) satisfies \(\mu \le \pi (\omega )\) for every \(\omega \in \Omega \), then \(X\) is \((1/\mu )\)-reasonable under \(P\); in particular,
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function of arity \(n\), let \(p\) be a natural number, and let \(P\) be a probability measure on the Boolean hypercube \(\{ 0,1\} ^n\) that assigns to each singleton \(\{ x\} \) the uniform weight \(2^{-n}\). Then the \(p\)-th moment of \(f\) under \(P\) equals the expectation of \(f^p\) under the uniform measure:
Let \(n\in \mathbb {N}\), let \(f\colon \{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function, and let \(\rho \in \mathbb {R}\). Write \(T_\rho \) for the noise operator with parameter \(\rho \), \(\langle \, \cdot \, ,\, \cdot \, \rangle \) for the \(L^2\) inner product under the uniform measure on \(\{ 0,1\} ^n\), and \(\hat f(S)\) for the Fourier–Walsh coefficient of \(f\) at frequency \(S\subseteq [n]\). Then
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(q \ge 1\) and let \(\rho \) be a real number with \(0 \le \rho \le 1\). Then for every Boolean function \(g \colon \{ 0,1\} ^1 \to \mathbb {R}\) on a single bit, the noise operator \(T_\rho \) satisfies
where both expectations are taken with respect to the uniform measure on \(\{ 0,1\} ^1\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix a natural number \(n\), exponents \(p,q\gt 1\) and a noise parameter \(\rho \) with \(0 \le \rho \le 1\), and write \(p' = p/(p-1)\) and \(q' = q/(q-1)\) for the Hölder conjugates. Then the noise operator \(T_\rho \) on Boolean functions \(f : \{ 0,1\} ^n \to \mathbb {R}\) satisfies the \((p \to q)\) bound
if and only if it satisfies the conjugate \((q' \to p')\) bound
where all norms are taken with respect to the uniform measure on \(\{ 0,1\} ^n\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(p,p'\gt 1\) be real Hölder conjugate exponents, so that \(1/p + 1/p' = 1\). Write \(\mathbb {E}\) for uniform expectation on \(\{ 0,1\} ^1\) and \(T_{\sqrt{p-1}}\) for the Fourier multiplier that sends the coefficient at a coordinate subset \(S\) to \((\sqrt{p-1})^{|S|}\) times that coefficient. Suppose that
holds for every one-bit Boolean function \(f\). Then for every one-bit Boolean function \(g\),
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(f:\{ 0,1\} ^{n+1}\to \mathbb {R}\) be a Boolean function, let \(\rho \in \mathbb {R}\), and let \(q\) be a nonnegative integer. Write \(A\) for the average of \(f\) over its last coordinate and \(D\) for the half-difference of \(f\) over its last coordinate, defined by \(A(x)=(f(x,0)+f(x,1))/2\) and \(D(x)=(f(x,0)-f(x,1))/2\) on \(\{ 0,1\} ^n\). Then the \(q\)-th moment of \(T_\rho f\) over the uniform measure on \(\{ 0,1\} ^{n+1}\) decomposes as
where on the right-hand side the noise operator acts on functions of \(n\) variables. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(n\) be a natural number. For \(\rho \in \mathbb {R}\) and \(x,y\in \{ 0,1\} ^n\), define
where \(\sigma (0)=1\) and \(\sigma (1)=-1\). For \(|\rho |\le 1\), this is the conditional transition probability from \(x\) to \(y\) under \(\rho \)-correlated noise; the joint probability of a correlated pair with uniform first point is \(2^{-n}K_\rho (x,y)\). The product defines a real-valued kernel for every real \(\rho \).
Let \(n\) be a natural number and let \(\rho \) be real with \(0\le \rho \le 1\). For every \(x,y\in \{ 0,1\} ^n\),
where \(\sigma (0)=1\) and \(\sigma (1)=-1\).
Fix a natural number \(n\), \(\rho \in \mathbb {R}\), and points \(x', y'\) in the Boolean hypercube \(\{ 0,1\} ^n\), and bits \(b, b'\). Write \((x', b)\) and \((y', b')\) for the points of \(\{ 0,1\} ^{n+1}\) obtained by appending \(b\) and \(b'\), respectively, as the final coordinate. Then the noise kernel of the extended pair factors as
where \(\chi (0)=1\) and \(\chi (1)=-1\).
Let \(n\) be a natural number, let \(0\le \rho \le 1\), and fix \(y\in \{ 0,1\} ^n\). The noise kernel \(K_\rho (x,y)=\prod _{i=1}^{n}\frac{1+\rho \, \sigma (x_i)\, \sigma (y_i)}{2}\), where \(\sigma (0)=1\) and \(\sigma (1)=-1\), satisfies
Thus \(K_\rho (\cdot ,y)\) is a probability distribution on the cube.
Let \(n\) be a natural number, let \(0\le \rho \le 1\), and fix \(x\in \{ 0,1\} ^n\). The noise kernel \(K_\rho (x,y)=\prod _{i=1}^{n}\frac{1+\rho \, \sigma (x_i)\, \sigma (y_i)}{2}\), where \(\sigma (0)=1\) and \(\sigma (1)=-1\), satisfies
Let \(n\) be a natural number, and let \(T_\rho \) denote the noise operator with parameter \(\rho \in [0,1]\) acting on Boolean functions \(f : \{ 0,1\} ^n \to \mathbb {R}\). Then for every real exponent \(s \ge 1\) and every point \(x \in \{ 0,1\} ^n\),
where \(K_\rho (x,y) = \prod _{i} \tfrac {1 + \rho \, \sigma (x_i)\, \sigma (y_i)}{2}\) is the noise kernel, with \(\sigma (0)=1\) and \(\sigma (1)=-1\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
For any real parameters \(\rho ,\sigma \) and any Boolean function \(f\colon \{ 0,1\} ^n\to \mathbb {R}\), applying the noise operator \(T_\sigma \) and then the noise operator \(T_\rho \) is the same as applying the single noise operator \(T_{\rho \sigma }\); that is,
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix \(\rho \in \mathbb {R}\), an arity \(n\), and a point \(x \in \{ 0,1\} ^n\). For every Boolean function \(g : \{ 0,1\} ^n \to \mathbb {R}\), the value of the noise operator \(T_\rho \) at \(x\) equals the kernel-weighted sum of \(g\), namely
where \(K_\rho (x,y) = \prod _{i} \tfrac {1 + \rho \, \sigma (x_i)\, \sigma (y_i)}{2}\) is the noise kernel and \(\sigma \) is the \(\pm 1\) sign encoding of a bit. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix \(\rho \in \mathbb {R}\) and a Boolean function \(f \colon \{ 0,1\} ^{n+1} \to \mathbb {R}\), and let \(Af\) and \(Df\) be the average and the half-difference of \(f\) over its last coordinate, each a Boolean function on \(\{ 0,1\} ^n\), with \(Af(x)=(f(x,0)+f(x,1))/2\) and \(Df(x)=(f(x,0)-f(x,1))/2\). For every \(x \in \{ 0,1\} ^n\) and every bit \(b\),
where \((x,b) \in \{ 0,1\} ^{n+1}\) is the point obtained by appending \(b\) as the last coordinate, \(\sigma (0)=1\), \(\sigma (1)=-1\), and \(T_\rho \) is the noise operator. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(n\) be a natural number, let \(p \ge 1\) and let \(f : \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a Boolean function on the \((n+1)\)-dimensional cube. For a point \(x' \in \{ 0,1\} ^n\), write \((x',0)\) and \((x',1)\) for the two extensions of \(x'\) obtained by appending a final bit. Then averaging over the last coordinate leaves the expected \(p\)-th power of \(\left\lvert f\right\rvert \) unchanged:
where both expectations are taken with respect to the uniform measure.
Let \(n\) be a natural number, let \(p \gt 0\) and let \(f : \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a Boolean function of arity \(n+1\). Writing \((x', b) \in \{ 0,1\} ^n \times \{ 0,1\} \) for the point of \(\{ 0,1\} ^{n+1}\) whose first \(n\) coordinates are \(x'\) and whose last coordinate is \(b\), one has
where both expectations are taken with respect to the uniform measure on the respective hypercubes.
Fix a real exponent \(q \ge 2\), and let \(g : \{ 0,1\} ^1 \to \mathbb {R}\) be a Boolean function of arity one. Write \(\lVert h\rVert _p = \left(\mathbb {E}\, |h|^{p}\right)^{1/p}\) for the \(L^p\) norm taken with respect to the uniform measure, and let \(T_\rho \) denote the noise operator. Then, with noise parameter \(\rho = 1/\sqrt{q-1}\),
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix a natural number \(n\) and \(p \gt 0\), let \(f : \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a Boolean function, and let \(x' \in \{ 0,1\} ^n\). Writing \((x', b) \in \{ 0,1\} ^{n+1}\) for the point obtained by appending a bit \(b \in \{ 0,1\} \) to \(x'\), and taking the expectation over the single free bit under the uniform measure, one has
Let \(f\colon \{ 0,1\} ^1\to \mathbb {R}\) be a Boolean function on a single bit, and let \(p\) be a real number with \(1\le p\le 2\). Suppose \(\rho \ge 0\) satisfies \(\rho ^2\le p-1\). Writing \(T_\rho \) for the noise operator and \(\mathbb {E}\) for expectation under the uniform measure on \(\{ 0,1\} ^1\),
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix a natural number \(n\), a real number \(\rho \), Boolean functions \(f, g : \{ 0,1\} ^{n+1} \to \mathbb {R}\), and points \(x', y' \in \{ 0,1\} ^n\). For a bit \(b\), write \((x', b)\) for the point of \(\{ 0,1\} ^{n+1}\) obtained by appending \(b\) as the last coordinate of \(x'\), and let \(\sigma \) denote the sign encoding with \(\sigma (0)=1\) and \(\sigma (1)=-1\). Then
where \(F, G : \{ 0,1\} ^1 \to \mathbb {R}\) are the one-bit slices \(F(t) = f(x', t_0)\) and \(G(t) = g(y', t_0)\), and the inner product and the noise operator \(T_\rho \) are taken on the one-dimensional cube \(\{ 0,1\} ^1\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(f:\{ 0,1\} ^1\to \mathbb {R}\) be a Boolean function of arity \(1\), and let \(x_0\in \{ 0,1\} ^1\) be the input whose single coordinate is \(\mathit{false}\) (that is, the sign encoding of that coordinate is \(+1\)). Then the value of \(f\) at this input equals the sum of its degree-zero and degree-one Fourier–Walsh coefficients:
Let \(f\colon \{ 0,1\} ^1\to \mathbb {R}\) be a function of a single bit, and write \(\hat f(\emptyset )\) and \(\hat f(\{ 0\} )\) for its two Fourier–Walsh coefficients, where \(\{ 0\} \) is the singleton frequency consisting of the sole coordinate. If \(x_1\) is the input whose single coordinate is \(1\) (the bit value corresponding to \(-1\) under the sign encoding), its value is the difference of these two coefficients,
Fix a natural number \(n\). Let \(1 \le p \le q\) with \(q \ge 2\), and let \(\rho \) satisfy \(0 \le \rho \le 1\) together with \(\rho \le \sqrt{(p-1)/(q-1)}\). Then the one-function hypercontractive inequality
holds if and only if the dual two-function inequality
holds, where \(T_\rho \) is the noise operator and \(\mathbb {E}\) and \(\langle \, \cdot \, ,\cdot \, \rangle \) are the expectation and \(L^2\) inner product with respect to the uniform measure on \(\{ 0,1\} ^n\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(Z\) be a measurable real-valued random variable on a probability space that is almost surely nonnegative and has finite first and second moments \(\mathbb {E}[Z]\) and \(\mathbb {E}[Z^2]\), with \(\mathbb {E}[Z] \gt 0\). Then for every \(\theta \) with \(0 \le \theta \le 1\),
Let \(q\) be a natural number and let \(f:\{ 0,1\} ^{n+1}\to \mathbb {R}\) be a Boolean function. Write \(g\) for the average of \(f\) over its last coordinate and \(h\) for the half-difference of \(f\) over its last coordinate, defined by \(g(x)=(f(x,0)+f(x,1))/2\) and \(h(x)=(f(x,0)-f(x,1))/2\) on \(\{ 0,1\} ^n\). Then the \(q\)-th moment of \(f\) under the uniform measure on \(\{ 0,1\} ^{n+1}\) decomposes as
where the expectation on the right is taken over the uniform measure on \(\{ 0,1\} ^n\).
Let \(n\) be a natural number, let \(f\colon \{ 0,1\} ^{n+1}\to \mathbb {R}\), and let \(b\in \{ 0,1\} \). The last-coordinate restriction of \(f\) at \(b\) is the function \(f_b\colon \{ 0,1\} ^n\to \mathbb {R}\) defined by \(f_b(x)=f(x,b)\).
Let \(f\colon \{ 0,1\} ^{n+1}\to \mathbb {R}\), and let \(x\) be a point of the \(n\)-dimensional Boolean hypercube. Then the restriction of \(f\) obtained by fixing its last coordinate to \(\mathrm{false}\) equals the sum of the average and the half-difference of \(f\) over that coordinate:
Let \(f : \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a Boolean function on \(n+1\) variables. For a bit \(b\), write \(f_b : \{ 0,1\} ^n \to \mathbb {R}\) for the restriction \(x \mapsto f(x,b)\) that fixes the last coordinate to \(b\), and let its average and half-difference over the last coordinate be the functions \(\tfrac 12(f_0 + f_1)\) and \(\tfrac 12(f_0 - f_1)\). Then for every point \(x \in \{ 0,1\} ^n\), the restriction fixing the last coordinate to \(\mathrm{true}\) equals the average minus the half-difference at that point:
Let \(f : \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a Boolean function on \(n+1\) variables. Define its average and half-difference over the last coordinate as the functions on \(n\) variables
Then, with the expectation on the left taken under the uniform measure on \(\{ 0,1\} ^{n+1}\) and those on the right under the uniform measure on \(\{ 0,1\} ^n\), the second moment of \(f\) decomposes as
For every real number \(x\), the product of its sign and itself equals its absolute value: \(\operatorname {sign}(x)\cdot x = \left\lvert x\right\rvert \), where \(\operatorname {sign}(x)\) is \(1\) if \(x\gt 0\), \(-1\) if \(x\lt 0\), and \(0\) if \(x=0\).
Let \(n\) be a natural number and let \(\varphi \colon \{ 0,1\} ^{n+1} \to \mathbb {R}\) be a real-valued function on the Boolean hypercube of dimension \(n+1\). Then the sum of \(\varphi \) over all points of the cube splits according to the value of the final coordinate:
where \((y,b)\) denotes the point of \(\{ 0,1\} ^{n+1}\) obtained by appending the bit \(b\) to \(y \in \{ 0,1\} ^{n}\) as its last coordinate.
For every natural number \(n\), every real number \(\rho \) and all points \(x, y\) of the Boolean hypercube \(\{ 0,1\} ^n\),
where the sum ranges over all subsets \(S\) of \([n]=\{ 1,\dots ,n\} \), \(\chi _S\) is the Walsh–Fourier character indexed by \(S\), and \(\sigma \) is the sign encoding sending \(0\mapsto 1\) and \(1\mapsto -1\).
Let \(n\) be a natural number, let \(s \ge 1\) and \(\rho \in [0,1]\), and let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function. Then, with expectations taken under the uniform measure on the cube, the noise operator \(T_\rho \) satisfies
For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(n\) be a natural number, let \(p,q\in \mathbb {R}\) satisfy \(p,q\ge 1\) and let \(\rho \) satisfy \(0 \le \rho \le 1\), and write \(T_\rho \) for the noise operator with parameter \(\rho \). Suppose the two-function bound
holds for every pair of functions \(f, g\) on the one-dimensional cube \(\{ 0,1\} ^1\), and also for every pair of functions on the \(n\)-dimensional cube \(\{ 0,1\} ^n\). Then the same bound holds for every pair of functions \(f, g : \{ 0,1\} ^{n+1} \to \mathbb {R}\). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(p, q \ge 1\) and let \(\rho \in \mathbb {R}\) be arbitrary. For any two Boolean functions \(f, g : \{ 0,1\} ^0 \to \mathbb {R}\) on the zero-dimensional cube, the \(L^2\) inner product of \(f\) with the noise-operator image \(T_\rho g\) satisfies
where \(\mathbb {E}\) denotes expectation under the uniform measure. For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(p,q,b\in \mathbb {R}\) satisfy \(1\le p\le q\le 2\) and \(0\le b\le 1\). Set \(\rho =\sqrt{(p-1)/(q-1)}\) when \(q\gt 1\) and \(\rho =0\) when \(q=1\). Then
Let \(n\) be a natural number and let \(x\) be a point of the Boolean hypercube \(\{ 0,1\} ^n\). Under the canonical uniform probability measure on \(\{ 0,1\} ^n\), the measure of the singleton \(\{ x\} \), taken as a real number, equals the uniform weight \(2^{-n}\).
For every natural number \(n\), the uniform weight \(2^{-n}\) assigned to each point of \(\{ 0,1\} ^n\) satisfies
so that the uniform weight on the cube of one higher dimension is half the uniform weight on the original cube.
Let \(p, q \in \mathbb {R}\) satisfy \(1 \le p \le 2\) and \(1 \le q \le 2\), and set \(\rho = \sqrt{(p-1)(q-1)}\). Then for every \(n \in \mathbb {N}\) and all Boolean functions \(f, g : \{ 0,1\} ^n \to \mathbb {R}\),
where \(\langle \cdot ,\cdot \rangle \) is the \(L^2\) inner product under the uniform measure on \(\{ 0,1\} ^n\), \(\mathbb {E}\) is the corresponding expectation, and \(T_\rho \) is the noise operator with parameter \(\rho \). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Let \(1 \le p \le 2\) and \(1 \le q \le 2\), and set \(\rho = \sqrt{(p-1)(q-1)}\). Then for any two Boolean functions \(f, g : \{ 0,1\} ^1 \to \mathbb {R}\) on a single bit,
where \(T_\rho \) is the noise operator, \(\mathbb {E}[\cdot ]\) is the expectation, and \(\left\langle \cdot ,\cdot \right\rangle \) is the \(L^2\) inner product, all taken with respect to the uniform measure on \(\{ 0,1\} \). For \(d\in \mathbb N\), \(h:\{ 0,1\} ^d\to \mathbb {R}\), and \(\tau \in \mathbb {R}\), the spectral noise operator is \(T_\tau h(x)=\sum _{S\subseteq [d]}\tau ^{|S|}\widehat h(S)\chi _S(x)\), where \(\chi _S(x)=(-1)^{\sum _{i\in S}x_i}\) and \(\widehat h(S)=\mathbb {E}[h\chi _S]\) under the uniform measure.
Fix a natural number \(n\), a real number \(\rho \), and a function \(F \colon \{ 0,1\} ^{n+1} \times \{ 0,1\} ^{n+1} \to \mathbb {R}\) on pairs of points of the Boolean hypercube of dimension \(n+1\). Write \(\chi (0)=1\) and \(\chi (1)=-1\) for the sign encoding of a bit, \(K_\rho (x,y)=\prod _i\frac{1+\rho \chi (x_i)\chi (y_i)}{2}\) for the noise kernel, and \((x',b)\) for the point of \(\{ 0,1\} ^{n+1}\) obtained by appending the bit \(b\) to \(x' \in \{ 0,1\} ^n\). Then
where \(x,y\) range over \(\{ 0,1\} ^{n+1}\), the points \(x',y'\) range over \(\{ 0,1\} ^{n}\), and the bits \(b,b'\) range over \(\{ 0,1\} \).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function and fix a coordinate \(i\in [n]\). Then the influence of \(i\) on \(f\) equals the sum of the squared Fourier–Walsh coefficients over all frequencies containing \(i\):
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function taking values in \(\{ -1, 1\} \), and let \(i\) be one of its \(n\) coordinates. Then the influence of coordinate \(i\) on \(f\) satisfies
where \(x^i\) denotes the point obtained from \(x\) by flipping its \(i\)-th coordinate and the expectation is taken under the uniform measure on \(\{ 0,1\} ^n\).
Let \(n\) be a natural number and let \(f, g, h : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions. With respect to the \(L^2\) inner product under the uniform measure, \(\langle f, g\rangle = 2^{-n}\sum _{x\in \{ 0,1\} ^n} f(x)\, g(x)\), one has
where \(f + g\) denotes the pointwise sum.
Let \(n\) be a natural number and let \(S \subseteq [n]\). Under the uniform measure on the Boolean hypercube \(\{ 0,1\} ^n\), the Walsh–Fourier character \(\chi _S\) satisfies
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a \(\pm 1\)-valued Boolean function, so that \(f(x) \in \{ -1, 1\} \) for every \(x \in \{ 0,1\} ^n\). Then its \(L^2\) inner product with itself, taken with respect to the uniform measure on the cube, equals \(1\):
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function on the hypercube of dimension \(n\ge 1\). Then some coordinate carries at least the average influence: there exists an index \(i\) with
where \(I[f]=\sum _{j=1}^{n}\mathrm{Inf}_j[f]\) is the total influence.
Let \(P\) be a probability measure on the Boolean hypercube \(\{ 0,1\} ^n\) that assigns every singleton the uniform weight \(2^{-n}\). Then for any Boolean function \(f : \{ 0,1\} ^n \to \mathbb {R}\) and any \(p \in \mathbb {N}\), the \(p\)-th moment of \(f\) under \(P\) equals the uniform expectation of \(f^p\):
The noise operator \(T_\rho \) with parameter \(\rho \in [-1,1]\), defined via the Fourier domain by
Probabilistically, \(T_\rho f(x) = \mathbb {E}_y[f(y)]\) where each bit of \(y\) agrees with \(x_i\) with probability \(\tfrac {1+\rho }{2}\) and is flipped with probability \(\tfrac {1-\rho }{2}\), independently.
Let \(\rho \in \mathbb {R}\), let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function, and let \(S\subseteq [n]\). Then the Fourier–Walsh coefficient of the noise-operator image \(T_\rho f\) at frequency \(S\) is
Fix a dimension \(n\) and let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions. For any real number \(\rho \), the noise operator \(T_\rho \) is self-adjoint with respect to the \(L^2\) inner product induced by the uniform measure on \(\{ 0,1\} ^n\):
For every \(n\), the parity function \(\mathrm{parity}_n(x) = \prod _{i=1}^{n}\sigma (x_i)\) on the Boolean hypercube \(\{ 0,1\} ^n\) coincides, as a Boolean function, with the Walsh–Fourier character \(\chi _{[n]}\) indexed by the full set \([n] = \{ 1,\dots ,n\} \); that is, \(\mathrm{parity}_n = \chi _{[n]}\).
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function, and equip \(\{ 0,1\} ^n\) with the uniform probability measure. Then the \(L^2\) inner product of \(f\) with itself equals the sum of the squares of its Fourier–Walsh coefficients over all subsets \(S \subseteq [n]\):
Let \(f \colon \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function that is \(\pm 1\)-valued, meaning \(f(x) \in \{ -1,1\} \) for every \(x \in \{ 0,1\} ^n\). Then its Fourier–Walsh coefficients satisfy
the sum ranging over all subsets \(S\) of \([n]\).
Let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions. Then their \(L^2\) inner product with respect to the uniform measure equals the sum, over all subsets \(S \subseteq [n]\), of the products of their corresponding Fourier–Walsh coefficients:
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function and let \(\rho \in \mathbb {R}\). Then the correlation of \(f\) with its noised version under the uniform measure is given by
where \(\langle \cdot ,\cdot \rangle \) is the \(L^2\) inner product, \(T_\rho \) is the noise operator, and \(\hat f(S)\) denotes the Fourier–Walsh coefficient of \(f\) at \(S\).
Let \(\sigma \colon \mathrm{Bool}\to \{ -1,1\} \) be the sign encoding, sending the value *false* to \(1\) and the value *true* to \(-1\). Summing \(\sigma \) over the two Boolean values yields
Let \(n\) be a natural number and let \(S \subseteq [n]\). Summing the Walsh–Fourier character \(\chi _S\) over the entire Boolean hypercube \(\{ 0,1\} ^n\) gives
Let \(x, y \in \{ 0,1\} ^n\) be points of the Boolean hypercube. Then the sum of the products of Walsh–Fourier characters over all subsets of \([n]\) satisfies
For any real numbers \(c_0,\dots ,c_{n-1}\), indexed by \(i \in \mathrm{Fin}\, n\),
where the sum ranges over all subsets \(S\) of the index set (with the empty product equal to \(1\)).
For every Boolean function \(f\colon \{ 0,1\} ^n\to \mathbb {R}\), the total influence of \(f\) is recovered from its Fourier–Walsh spectrum by
where the sum ranges over all subsets \(S\) of \([n]=\{ 1,\dots ,n\} \), \(\left\lvert S\right\rvert \) is the cardinality of \(S\), and \(\hat f(S)\) is the Fourier–Walsh coefficient of \(f\) at frequency \(S\).
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function of arity \(n\). Then at every point \(x\in \{ 0,1\} ^n\),
where the sum ranges over all subsets \(S\) of \([n]\), \(\chi _S\) is the Walsh character of \(S\), and \(\hat f(S)=\langle f,\chi _S\rangle \) is the Fourier–Walsh coefficient of \(f\) at \(S\).
Under the embedding \(\mathrm{BoolToPM1} : \mathrm{Bool} \to \mathbb {R}\) that sends \(\mathtt{false}\) to \(1\) and \(\mathtt{true}\) to \(-1\), the average of its two values is zero:
Let \(\mathrm{BoolToPM1} : \mathrm{Bool} \to \mathbb {R}\) be the Bool-to-\(\pm 1\) embedding, which sends \(\mathtt{false}\) to \(1\) and \(\mathtt{true}\) to \(-1\). Then for every Boolean value \(b\), negating \(b\) flips the sign of its image:
For every Boolean value \(b\), the image \(\mathrm{BoolToPM1}(b)\) under the map sending \(\mathtt{false}\) to \(1\) and \(\mathtt{true}\) to \(-1\) satisfies \(\mathrm{BoolToPM1}(b)\cdot \mathrm{BoolToPM1}(b) = 1\); that is, this \(\pm 1\)-valued quantity squares to \(1\).
Let \(\mathrm{BoolToPM1} : \mathrm{Bool} \to \mathbb {R}\) be the map sending \(\mathtt{false}\) to \(1\) and \(\mathtt{true}\) to \(-1\). Then for all Booleans \(a\) and \(b\),
so the exclusive-or \(a \oplus b\) is carried to the ordinary product in \(\{ -1,1\} \subset \mathbb {R}\).
Let \(n\) be a natural number. The Walsh character \(\chi _\emptyset \) associated to the empty subset of \([n]\) is the constant function equal to \(1\) on the Boolean hypercube; that is, \(\chi _\emptyset (x) = 1\) for every \(x \in \{ 0,1\} ^n\).
Let \(n\) be a natural number and let \(S \subseteq [n]\). Evaluating the Walsh character \(\chi _S : \{ 0,1\} ^n \to \mathbb {R}\) at the all-zeros point \(\mathbf{0}\) (the vector whose every coordinate is \(\mathtt{false}\)) gives
For any two subsets \(S, T \subseteq [n]\) and any point \(x \in \{ 0,1\} ^n\),
where \(\chi _S\) denotes the Walsh–Fourier character of \(S\) and \(S \mathbin {\triangle } T\) is the symmetric difference of \(S\) and \(T\).
For every \(n \in \mathbb {N}\), the Walsh character \(\chi _\emptyset \) associated to the empty subset of \([n]\) has uniform expectation equal to one:
Let \(n\) be a natural number and let \(S \subseteq [n]\) be a nonempty set of coordinates. Then the Walsh–Fourier character \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\) has vanishing uniform expectation over the Boolean hypercube \(\{ 0,1\} ^n\):
Let \(n\) be a natural number and let \(S \subseteq [n]\) be a nonempty set of coordinates. Then the uniform expectation of the Walsh character \(\chi _S\) over the Boolean hypercube factorises as a product over the coordinates in \(S\):
where \(\mathrm{BoolToPM1}\) sends \(\mathtt{false}\) to \(1\) and \(\mathtt{true}\) to \(-1\).
Let \(n\) be a natural number and let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function of arity \(n\). Then the two forms of the uniform expectation of \(f\) coincide, namely
Let \(g : \{ 0,1\} ^n \to \mathbb {R}\) be given coordinatewise, so that for each index \(i \in \{ 0,\dots ,n-1\} \) we have a function \(g_i : \{ 0,1\} \to \mathbb {R}\), and let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be the product function \(f(x) = \prod _{i} g_i(x_i)\). Then its uniform expectation over the Boolean hypercube factorises as
Let \(n\) be a natural number, and let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions. Then for every subset \(S \subseteq [n]\), the Fourier coefficient of the convolution \(f * g\) at \(S\) equals the product of the Fourier coefficients of \(f\) and \(g\) at \(S\):
Let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions, let \(S \subseteq [n]\), and write \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\) for the associated Walsh character, where \(\oplus \) denotes componentwise XOR. Suppose the order of summation may be interchanged in the double sum
Then for every \(y \in \{ 0,1\} ^n\), translating the argument of \(g\) by \(y\) multiplies the character-weighted sum by the factor \(\chi _S(y)\):
Let \(n\) be a natural number, let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions, and let \(S \subseteq [n]\) with associated Walsh character \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\). Writing \(x \oplus y\) for the componentwise XOR of \(x, y \in \{ 0,1\} ^n\), the following identity holds:
Let \(S \subseteq [n]\), let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions, and suppose the interchange-of-summation identity
holds, where all sums range over \(\{ 0,1\} ^n\) and \(\oplus \) is componentwise XOR. Then for all \(x, y \in \{ 0,1\} ^n\),
where \(\chi _S\) is the Walsh–Fourier character of \(S\); that is, \(\chi _S\) is a character of the group \((\mathbb {F}_2)^n\).
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function of arity \(n\), and let \(S \subseteq [n]\). Then the two definitions of the Fourier–Walsh coefficient \(\hat f(S)\) coincide: both are equal to the \(L^2\) inner product of \(f\) with the Walsh character \(\chi _S\),
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function and let \(x \in \{ 0,1\} ^n\). Then \(f\) agrees pointwise with the sum of its Walsh characters weighted by its Fourier coefficients:
where the sum ranges over all subsets \(S\) of \([n]\), \(\hat f(S) = \langle f, \chi _S\rangle \) is the Fourier coefficient of \(f\) at \(S\), and \(\chi _S\) is the Walsh character of \(S\).
Fix \(n \in \mathbb {N}\), and let \(S, T \subseteq [n]\) be two distinct subsets. Then the Walsh–Fourier characters \(\chi _S\) and \(\chi _T\) are orthogonal with respect to the uniform inner product on \(\{ 0,1\} ^n\):
Let \(n\) be a natural number and let \(S \subseteq [n]\). Then the Walsh–Fourier character \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\) satisfies
where the inner product is the uniform expectation of the pointwise product.
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a real-valued function on the Boolean hypercube, and for each subset \(S \subseteq [n]\) let \(\hat f(S) = \langle f, \chi _S\rangle \) be its Fourier coefficient, where the inner product and expectation are taken with respect to the uniform measure. Then the sum of the squared Fourier coefficients equals the squared \(L^2\) norm of \(f\):
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a real-valued function on the Boolean hypercube, and equip such functions with the uniform-measure inner product \(\langle g, h \rangle = 2^{-n}\sum _{x \in \{ 0,1\} ^n} g(x)\, h(x)\). Then the squared \(L^2\) norm of \(f\) equals the inner product of \(f\) with itself:
The partial inverse \(\mathrm{PM1ToBool?} : \mathbb {R} \to \mathrm{Option}\, \mathrm{Bool}\) returns \(\mathtt{some}\, \mathtt{true}\) on \(-1\), \(\mathtt{some}\, \mathtt{false}\) on \(1\), and \(\mathtt{none}\) on all other real numbers.
Fix \(n \in \mathbb {N}\), and for each subset \(S \subseteq [n]\) let \(\chi _S \colon \{ 0,1\} ^n \to \mathbb {R}\) denote the Walsh–Fourier character \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\). Writing \(\mathbf{0}\) for the all-zeros point of \(\{ 0,1\} ^n\), summing these characters evaluated at \(\mathbf{0}\) over all \(2^n\) subsets of \([n]\) gives
the number of points of the Boolean hypercube.
Fix \(n \ge 0\), and for each subset \(S \subseteq [n]\) let \(\chi _S : \{ 0,1\} ^n \to \mathbb {R}\) be the Walsh character \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\). Writing \(\mathbf{0}\) for the all-zeros point of \(\{ 0,1\} ^n\), the map sending each \(S\) to \(\chi _S(\mathbf{0})\) is the constant function equal to \(1\):
Fix \(n\in \mathbb {N}\) and a point \(x\in \{ 0,1\} ^n\), and for each subset \(S\subseteq [n]\) let \(\chi _S(x)=\prod _{i\in S}(-1)^{x_i}\) be the Walsh–Fourier character. If \(x\) is not the all-zeros point, then
the sum ranging over all \(2^n\) subsets of \([n]\).
Let \(x \in \{ 0,1\} ^n\) be a point of the Boolean hypercube that differs from the all-zeros point \(\mathbf{0}\), and suppose \(i\) is a coordinate at which \(x_i = \mathtt{true}\). Then summing the Walsh–Fourier characters over all subsets \(S \subseteq [n]\), evaluated at \(x\), gives the same value as summing their negations:
Let \(n\) be a natural number and let \(M\) be a subset of the \(n\)-element index set \(\mathrm{Fin}\, n = \{ 0, 1, \dots , n-1\} \). Then the size of \(M\) and the size of its complement \(M^c = \mathrm{Fin}\, n \setminus M\) add up to \(n\):
Let \(G\) be a simple graph on a finite vertex set. Then the number of orientations of \(G\) is \(2^{\left\lvert E(G)\right\rvert }\), where \(\left\lvert E(G)\right\rvert \) denotes the number of edges of \(G\); that is, each edge may be directed in either of its two ways, and every such choice yields a distinct orientation.
For natural numbers \(n\) and \(t\), the number of Pauli strings of length \(n\) whose weight is at most \(t\) is
Here a Pauli string is a function \(p : \mathrm{Fin}\, n \to \{ I,X,Y,Z\} \), and its weight \(\mathrm{wt}(p)\) is the number of coordinates \(i\) with \(p(i) \neq I\); the sum counts, for each \(j \le t\), the \(\binom {n}{j}\) choices of a weight-\(j\) support together with the \(3^{j}\) assignments of a non-identity Pauli to each chosen coordinate.
Let \(n, m\) be positive integers, and let \(v, w\) be real numbers (thought of as the value of a \(3\)-XOR polynomial and that of its derived \(4\)-XOR polynomial). If
then
Let \(n, t\) be natural numbers with \(t + 1 \le n\), and let \(a, x\) be real numbers with \(a \ne x\). Then there is a positive constant \(c \gt 0\) for which the truncated Christoffel–Darboux kernel satisfies
where \(K_j^{(n)}\) denotes the \(j\)-th Krawtchouk polynomial.
Fix natural numbers \(t\) and \(n\) with \(t + 1 \le n\), and let \(a\) be a real number lying strictly to the left of every real zero of the Krawtchouk polynomial \(K_t^{(n)}\); that is, \(a \lt \xi \) whenever \(K_t^{(n)}(\xi ) = 0\). Then there is a real threshold \(\eta \) such that the truncated Christoffel–Darboux kernel satisfies \(\Phi _{n,t}(a, x) \le 0\) for every integer \(x\) with \(\eta \le x \le n\).
Let \((\Omega ,\mu )\) be a probability space and let \(k\) be a positive natural number. Suppose \(Y:\Omega \to \mathbb {R}\) is measurable and normally distributed with mean \(0\) and variance \(1/k\), that is, its law \(Y_*\mu \) equals \(N(0,1/k)\). Then for every real \(t\) with \(\left\lvert t\right\rvert \le k/4\), the function \(\omega \mapsto \exp \! \bigl(t\, (Y(\omega )^2 - 1/k)\bigr)\) is \(\mu \)-integrable, and the moment generating function of the centered variable \(Y^2 - 1/k\) satisfies
Let \(\mu \) be a probability measure on \(\Omega \), let \(k \ge 1\), and let \(Y_1,\dots ,Y_k : \Omega \to \mathbb {R}\) be independent random variables, each with law \(N(0, 1/k)\). Then for every \(\varepsilon \) with \(0 \lt \varepsilon \lt 1\),
Let \(D\) be a digraph on a finite vertex set \(V\), let \(f : V \to \mathbb {R}\) be a real-valued function on the vertices, and let \(m\) and \(n\) be natural numbers. Write \(\chi (D)\) for the chromatic number of the underlying simple graph of \(D\) (in which two vertices are joined whenever there is an arc between them in either direction). If \(\chi (D) \gt mn\), then at least one of the following holds: either \(D\) contains a directed path \(u_0 \to u_1 \to \cdots \to u_m\) on \(m+1\) distinct vertices along which \(f\) is nondecreasing, \(f(u_0) \le f(u_1) \le \cdots \le f(u_m)\); or \(D\) contains a directed path \(v_0 \to v_1 \to \cdots \to v_n\) on \(n+1\) distinct vertices along which \(f\) is strictly decreasing, \(f(v_0) \gt f(v_1) \gt \cdots \gt f(v_n)\).
(Chvátal and Lovász, 1974.) Let \(D\) be a loopless digraph on a finite vertex set \(V\); that is, no vertex has an arc to itself. Then \(D\) admits an independent set \(S \subseteq V\) — no two distinct vertices of \(S\) are joined by an arc of \(D\) in either direction — such that every vertex \(v \notin S\) is reachable from some vertex \(u \in S\) by a directed path of length one or two: either \(u \to v\) is an arc of \(D\), or there is a vertex \(w\) with arcs \(u \to w\) and \(w \to v\).
Let \(p\) be prime and let \(S \le V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) be a subspace that is isotropic for the symplectic form \(\omega \), meaning \(\omega (u,v)=0\) for all \(u,v \in S\). Then for every subset \(M \subseteq \{ 0,\dots ,n-1\} \), with complement \(M^c\),
where \(g(S,M) = \dim _{\mathbb {F}_p}\! \big(S^{\perp }_M\big) - \dim _{\mathbb {F}_p}\! \big(S_M\big)\).
Converts a CNF \(\psi \) to its De Morgan dual DNF by negating every literal, so each clause (a disjunction) becomes a term (a conjunction) of negated literals: \(\lnot \bigl(\bigwedge _i \bigvee _j l_{ij}\bigr) = \bigvee _i \bigl(\bigwedge _j \lnot l_{ij}\bigr)\).
Let \(\psi \) be a formula in conjunctive normal form over \(n\) Boolean variables — a conjunction of clauses, each clause a disjunction of literals — and let its De Morgan dual DNF be obtained by negating every literal of \(\psi \), so that each clause \(\bigvee _j l_{ij}\) becomes the term \(\bigwedge _j \lnot l_{ij}\). Then for every assignment \(x \colon \{ 0,\dots ,n-1\} \to \{ \mathrm{true},\mathrm{false}\} \), evaluating this dual DNF at \(x\) yields exactly the negation of the value of \(\psi \) at \(x\).
Let \(\psi \) be a CNF formula on \(n\) variables with the property that, within each of its clauses, any two literals having the same variable are in fact the same literal (so no variable occurs in a clause with both polarities, nor is any literal repeated). Then the De Morgan dual DNF obtained from \(\psi \) by negating every literal has the same property term by term: within each of its terms, any two literals sharing a variable are equal.
Let \(\psi \) be a CNF formula whose clauses are lists of literals, and suppose that every clause of \(\psi \) contains no repeated literal. Then every term of the De Morgan dual DNF obtained from \(\psi \) — the DNF formed by negating each literal of each clause — likewise contains no repeated literal.
Let \(\psi \) be a CNF formula over \(n\) variables, and let \(\psi ^{*}\) be its De Morgan dual DNF, obtained by negating every literal so that each clause becomes a term. Then \(\psi ^{*}\) and \(\psi \) have the same width; that is, the maximum term-length of \(\psi ^{*}\) equals the maximum clause-width of \(\psi \).
Let \(p\) be a prime and let \(S\) be a subspace of the symplectic space \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\). If \(S\) is isotropic and its logical dimension vanishes, \(k(S) = 0\), then the code distance vanishes as well, \(d(S) = 0\).
Fix a prime \(p\) and a natural number \(n\), and let \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) carry the symplectic form \(\omega \bigl((x,z),(x',z')\bigr) = \sum _{i} (x_i z'_i - z_i x'_i)\). For every \(\mathbb {F}_p\)-subspace \(S \le V\), the associated code distance \(d(S)\) — the least weight \(\mathrm{wt}(v) = |\mathrm{supp}(v)|\) attained by a vector \(v\) lying in the symplectic orthogonal complement \(S^{\perp _\omega }\) but not in \(S\) — satisfies \(d(S) \le n\).
For naturals \(n\) and \(d\), \(A(n,d)\) is the supremum of the cardinalities \(k\) for which there is a finite set \(C\) of binary codewords \(\mathrm{Fin}\, n \to \mathrm{Bool}\) with \(|C| = k\) and such that any two distinct \(x, y \in C\) differ in at least \(d\) coordinates.
Let \(C\) be a closed subspace of the \(n\)-qubit Hilbert space \(\mathcal{H}_n\), and let \(P_C : \mathcal{H}_n \to \mathcal{H}_n\) be the associated orthogonal projection onto \(C\). Then for every \(x \in \mathcal{H}_n\), the value \(P_C\, x\) is exactly the orthogonal projection of \(x\) onto \(C\).
Let \(C\) be a closed subspace of the \(n\)-qubit Hilbert space \(\mathcal{H}_n = \ell ^2(\{ 0,1\} ^n, \mathbb {C})\), and let \(P_C : \mathcal{H}_n \to \mathcal{H}_n\) be the orthogonal projection onto \(C\). Then \(P_C\) is idempotent: for every \(x \in \mathcal{H}_n\),
Let \(C\) be a subspace of the \(n\)-qubit Hilbert space \(\mathcal{H}_n\), and let \(P_C : \mathcal{H}_n \to \mathcal{H}_n\) be the orthogonal projection onto \(C\). Then for every \(x \in \mathcal{H}_n\), the projected vector \(P_C\, x\) belongs to \(C\).
Let \(n \ge 1\) and let \(d\) be an integer with \(1 \le d\) and \(2d \le n\). Suppose \(C\) is a finite set of binary words of length \(n\) such that any two distinct words of \(C\) are at Hamming distance at least \(d\), and every word of \(C\) has Hamming weight at most \(w\). If \(w \le J_2(n,d)\), where \(J_2(n,d) = \tfrac {1}{2}\bigl(n - \sqrt{n(n-2d)}\bigr)\) is the binary Johnson radius, then \(\left\lvert C\right\rvert \le 2n\).
Let \(n \ge 1\), \(d \ge 1\), and suppose \(2d \le n\). If \(C\) is an \((n,d,w)\)-admissible code — a finite set of binary words of length \(n\) in which any two distinct words are at Hamming distance at least \(d\) and every word has weight at most \(w\) — and the weight bound satisfies \(w \le J_2(n,d)\), where \(J_2(n,d) = \tfrac {1}{2}\bigl(n - \sqrt{n(n-2d)}\bigr)\), then \(\left\lvert C\right\rvert \le 2n\).
Let \(n \ge 1\), let \(d, w\) be natural numbers, and let \(C\) be a finite set of binary words of length \(n\) such that any two distinct words of \(C\) are at Hamming distance at least \(d\), and every word of \(C\) has weight at most \(w\). Suppose \(\alpha \ge 0\) is a real parameter for which the shifted \(\pm 1\) vector \(\hat{x}^\alpha \) is nonzero for every \(x \in C\), and for which
Then \(\left\lvert C\right\rvert \le 2n\).
Let \(x, y\) be binary words of length \(n\), and encode each as a \(\pm 1\) vector in \(\mathbb {R}^n\) by sending its \(i\)-th bit to \(-1\) when that bit is true and to \(+1\) when it is false. Then for every coordinate \(i\), the product of the \(i\)-th entries of these two embeddings equals \(1\) if \(x_i = y_i\) and \(-1\) if \(x_i \ne y_i\).
Let \(V\) be a finite-dimensional real inner product space, and let \(u \in V\) be a unit vector, so that \(\lVert u\rVert = 1\). Then the orthogonal complement of the line \(\mathrm{span}_{\mathbb {R}}\{ u\} \) has dimension one less than that of \(V\):
Let \(x\) be a binary word of length \(n\), and let \(s(x) \in \mathbb {R}^n\) denote its \(\pm 1\) embedding, the vector whose \(i\)-th coordinate is \(-1\) at each position where \(x_i\) is set and \(+1\) elsewhere. Writing \(\mathbf{1}\) for the all-ones vector and \(\mathrm{wt}(x)\) for the Hamming weight of \(x\) (the number of positions at which \(x\) is set), the \(\ell ^2\) inner product satisfies
For any two binary words \(x, y\) of length \(n\), let \(\phi (x), \phi (y) \in \mathbb {R}^n\) be their \(\pm 1\) embeddings, whose \(i\)-th coordinates are \(-1\) when the corresponding bit is true and \(+1\) otherwise. Then their Euclidean inner product is
where \(d(x,y)\) is the Hamming distance, the number of coordinates on which \(x\) and \(y\) differ.
Let \(V\) be a real inner product space, let \(u \in V\) be a unit vector, and let \(x, y \in V\) satisfy \(\langle x, u\rangle \le 0\), \(\langle y, u\rangle \le 0\), and \(\langle x, y\rangle \le 0\). Then the components of \(x\) and \(y\) orthogonal to \(u\) again have nonpositive inner product:
Let \(\alpha \in \mathbb {R}\), let \(x, y\) be binary words of length \(n\), and write \(\mathbf{1}\) for the all-ones vector in \(\mathbb {R}^n\) and \(s(x), s(y) \in \mathbb {R}^n\) for the \(\pm 1\) embeddings of \(x\) and \(y\) (with \(i\)-th coordinate \(-1\) where the corresponding bit is true and \(+1\) otherwise). Form the shifted vectors \(\hat{x}^\alpha = s(x) - \alpha \mathbf{1}\) and \(\hat{y}^\alpha = s(y) - \alpha \mathbf{1}\). Then their Euclidean inner product expands as
Let \(x\) and \(y\) be binary words of length \(n\), and let \(\alpha \ge 0\). Suppose \(d\) and \(w\) are natural numbers for which the Hamming distance between \(x\) and \(y\) is at least \(d\) and the Hamming weights of \(x\) and \(y\) are each at most \(w\). Writing \(\hat{x}^\alpha \) and \(\hat{y}^\alpha \) for the corresponding shifted vectors in \(\mathbb {R}^n\) — the \(\pm 1\) embedding of the word minus \(\alpha \) times the all-ones vector — their Euclidean inner product satisfies
Let \(n\), \(d\), and \(w\) be natural numbers with \(n \gt 0\) and \(2d \le n\), and write \(\alpha = \alpha (n,d)\) for the shift parameter and \(J_2(n,d)\) for the binary Johnson radius. If \(w \le J_2(n,d)\), then
Given a unit vector \(u\), this packages the normalized projection \(\lVert \mathrm{orthProj}(u,v)\rVert ^{-1}\mathrm{orthProj}(u,v)\) together with its membership proof as an element of the subspace \((\mathrm{span}_{\mathbb {R}}\{ u\} )^{\perp }\).
Let \(V\) be a real inner product space, let \(u \in V\) be a unit vector, and let \(v \in V\). Write \(w = v - \langle u, v\rangle \, u\) for the component of \(v\) orthogonal to \(u\). The normalized vector \(\lVert w\rVert ^{-1} w\), regarded as an element of the orthogonal complement \((\mathrm{span}_{\mathbb {R}}\{ u\} )^{\perp }\) of the line spanned by \(u\), has underlying vector in \(V\) equal to \(\lVert w\rVert ^{-1} w\).
Let \(u,v\) be unit vectors in a real inner product space, with \(v \neq u\) and \(v \neq -u\), and set \(w = v - \langle u,v\rangle \, u\), the projection of \(v\) onto the orthogonal complement of \(\mathrm{span}\{ u\} \). Then the rescaled vector \(\lVert w\rVert ^{-1}\, w\) has norm \(1\).
Let \(u,v,w\) be unit vectors in a real inner product space satisfying \(\langle v,u\rangle \le 0\), \(\langle w,u\rangle \le 0\), and \(\langle v,w\rangle \le 0\), and suppose that neither \(v\) nor \(w\) equals \(u\) or \(-u\) and that \(v \neq w\). For \(x\) in the space write \(\pi (x) = x - \langle u,x\rangle \, u\) for the projection of \(x\) onto the orthogonal complement of \(\mathrm{span}\{ u\} \). Then the two normalized projections \(\lVert \pi (v)\rVert ^{-1}\, \pi (v)\) and \(\lVert \pi (w)\rVert ^{-1}\, \pi (w)\) cannot be equal.
Let \(u,v,w\) be vectors in a real inner product space with \(\lVert u\rVert =1\), and suppose the three inner products \(\langle v,w\rangle \), \(\langle v,u\rangle \), and \(\langle w,u\rangle \) are all nonpositive. For a vector \(x\), write \(x^{\perp }=x-\langle u,x\rangle \, u\) for its component orthogonal to \(u\), and assume that both \(v^{\perp }\) and \(w^{\perp }\) are nonzero. Then their normalizations still have nonpositive inner product:
Let \(V\) be a real inner product space, and let \(u,v,w \in V\) with \(\lVert u\rVert = 1\). Suppose the three pairwise inner products \(\langle v,w\rangle \), \(\langle v,u\rangle \), and \(\langle w,u\rangle \) are all nonpositive. For a vector \(x \in V\), write \(x - \langle u,x\rangle \, u\) for the projection of \(x\) onto the orthogonal complement of \(\mathrm{span}\{ u\} \). Then the projections of \(v\) and \(w\) again have nonpositive inner product:
Let \(V\) be a real inner product space, and let \(u,v \in V\) with \(\lVert u\rVert = 1\). Then the vector \(v - \langle u,v\rangle \, u\) lies in the orthogonal complement \((\mathrm{span}_{\mathbb {R}}\{ u\} )^{\perp }\) of the line spanned by \(u\).
Let \(u\) and \(v\) be unit vectors in a real inner product space, and suppose that \(v \neq u\) and \(v \neq -u\). Then the projection of \(v\) onto the orthogonal complement of \(\mathrm{span}\{ u\} \), namely \(v - \langle u,v\rangle \, u\), is nonzero.
Let \(x\) be a binary word of length \(n\), and let the \(\pm 1\) embedding send \(x\) to the vector in \(\mathbb {R}^n\) whose \(i\)-th coordinate is \(-1\) when \(x_i\) is true and \(+1\) when \(x_i\) is false. If the \(i\)-th entry of \(x\) is false, then the \(i\)-th coordinate of the embedded vector equals \(1\).
Let \(x\) be a binary word of length \(n\), and let \(s(x) \in \mathbb {R}^n\) denote its \(\pm 1\) embedding, whose \(i\)-th coordinate is \(-1\) when \(x_i\) is true and \(+1\) when \(x_i\) is false. Then for every index \(i\) with \(x_i\) true, the \(i\)-th coordinate of \(s(x)\) equals \(-1\).
Let \(V\) be a real inner product space, and let \(u \in V\) be a unit vector. Let \(S \subseteq V\) be a set of unit vectors each of which makes a non-positive inner product with \(u\); that is, \(\lVert x\rVert = 1\) and \(\langle x, u\rangle \le 0\) for every \(x \in S\). Then the map \(x \mapsto x - \langle x, u\rangle \, u\), which subtracts from each vector its component along \(u\), is injective on \(S\).
Let \(V\) be a real inner product space, and let \(u, x \in V\) be unit vectors, \(\lVert u\rVert = 1\) and \(\lVert x\rVert = 1\), with \(\langle x, u\rangle \le 0\) and \(x \neq u\), \(x \neq -u\). Then the component of \(x\) orthogonal to \(u\) is nonzero, that is, \(x - \langle x, u\rangle \, u \neq 0\).
Let \(V\) be a real inner product space, let \(u \in V\) be a unit vector, so \(\lVert u\rVert = 1\), and let \(x \in V\) be arbitrary. Then
Let \(V\) be a finite-dimensional real inner product space, and let \(S\) be a finite set of vectors in \(V\), each of norm \(1\), such that any two distinct vectors of \(S\) have non-positive inner product. Then \(\left\lvert S\right\rvert \le 2 \dim _{\mathbb {R}} V\).
Let \(S\) be a finite set of vectors in the Euclidean space \(\mathbb {R}^n\), each of norm \(1\), such that any two distinct vectors \(u, v \in S\) satisfy \(\langle u, v\rangle \le 0\). Then \(\left\lvert S\right\rvert \le 2n\).
Let \(n \ge 1\) and let \(\alpha \) be a real number with \(0 \le \alpha \lt 1\). Then for every binary word \(x\) of length \(n\), the shifted vector \(\hat{x}^\alpha \in \mathbb {R}^n\) — obtained from the \(\pm 1\) embedding of \(x\) by subtracting \(\alpha \) times the all-ones vector — is nonzero.
For all \(n, t \in \mathbb {N}\), write \(V_q(n,t) = \sum _{i=0}^{t} \binom {n}{i}(q-1)^i\) for the ball volume at radius \(t\) of a length-\(n\) code over an alphabet of size \(q\). Over a binary alphabet, i.e. taking \(q = 2\), this reduces to the partial sum of binomial coefficients
Let \(x\) and \(y\) be \(n\)-bit strings. Then the number of coordinates on which \(x\) and \(y\) agree, together with their Hamming distance \(d(x, y)\) — the number of coordinates on which they differ — sums to the length:
Let \(x, y : \mathrm{Fin}\, n \to \mathrm{Bool}\) be two \(n\)-bit strings, with agreement count \(\mathrm{agr}(x,y)\) the number of coordinates on which they coincide and \(d_H(x,y)\) their Hamming distance (the number of coordinates on which they differ). Then, as an identity of integers,
Given two \(n\)-bit strings \(x, y : \mathrm{Fin}\, n \to \mathrm{Bool}\), their signed inner product is the integer
Each coordinate where \(x\) and \(y\) agree contributes \(+1\), and each disagreeing coordinate contributes \(-1\).
Let \(x, y \colon \{ 0, \dots , n-1\} \to \{ 0,1\} \) be two \(n\)-bit strings, and write \(\langle x, y\rangle \) for their signed inner product, the integer \(\sum _{i} \sigma _i\) where \(\sigma _i = +1\) when \(x_i = y_i\) and \(\sigma _i = -1\) otherwise. For a natural number \(a\), let \(x^{(a)}\) and \(y^{(a)}\) denote the \(an\)-bit strings obtained by concatenating \(a\) consecutive copies of \(x\) and of \(y\), respectively. Then
Let \(x_1, y_1 : \mathrm{Fin}\, m \to \mathrm{Bool}\) be \(m\)-bit strings and \(x_2, y_2 : \mathrm{Fin}\, n \to \mathrm{Bool}\) be \(n\)-bit strings, and write \(x_1 \frown x_2\) for the length-\((m+n)\) string obtained by concatenation, whose first \(m\) coordinates are those of \(x_1\) and whose last \(n\) coordinates are those of \(x_2\). Writing \(\langle x, y\rangle \) for the signed inner product of two equal-length bit strings — the number of coordinates on which they agree minus the number on which they disagree — the signed inner product of the two concatenations splits as the sum of the block-wise signed inner products:
Let \(m\) and \(n\) be natural numbers with \(m = n\), and let \(x, y \colon \mathrm{Fin}\, n \to \mathrm{Bool}\) be two \(n\)-bit strings, writing \(\langle x, y\rangle \) for their signed inner product. The equality \(m = n\) gives a bijection \(\sigma \colon \mathrm{Fin}\, m \to \mathrm{Fin}\, n\), and precomposing with it turns \(x\) and \(y\) into the \(m\)-bit strings \(x \circ \sigma \) and \(y \circ \sigma \). Then reindexing both strings in this way leaves the signed inner product unchanged:
Let \(x\) and \(y\) be two \(n\)-bit strings. Their signed inner product \(\langle x, y\rangle \) is the integer \(\sum _{i} \varepsilon _i\), where \(\varepsilon _i = +1\) at each coordinate \(i\) on which \(x_i = y_i\) and \(\varepsilon _i = -1\) at each coordinate on which \(x_i \neq y_i\). Writing \(a(x,y)\) for the agreement count of \(x\) and \(y\), the number of coordinates on which they agree, and \(d_H(x,y)\) for the Hamming distance between them, the number of coordinates on which they differ, one has
an equality of integers in which the two counts are regarded as integers.
Let \(x, y\) be two \(n\)-bit strings, that is, functions from \(\{ 0, 1, \dots , n-1\} \) to \(\{ 0,1\} \), and let \(d(x,y)\) denote their Hamming distance, the number of coordinates on which they disagree. Then their signed inner product — the sum over all coordinates of \(+1\) where \(x\) and \(y\) agree and \(-1\) where they disagree — satisfies
For all Boolean values \(a\) and \(b\), the product of their signs satisfies
where \(\mathrm{sign}(\cdot )\) sends \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\), and \(\mathbf{1}[a \ne b]\) is \(1\) when \(a\) and \(b\) differ and \(0\) when they agree. Equivalently, this product equals \(1\) when the two bits are equal and \(-1\) when they disagree.
Let \(\alpha \) be a type, let \(s\) be a finite set of indices from \(\alpha \), and let \(f : \alpha \to \{ \mathrm{false}, \mathrm{true}\} \) be a Boolean-valued function. Write \(\sigma \) for the Boolean sign, so that \(\sigma (\mathrm{false}) = 1\) and \(\sigma (\mathrm{true}) = -1\), and let \(\bigoplus _{i \in s} f(i)\) denote the sum of the values \(f(i)\) taken in the Booleans under exclusive-or, so that this sum equals \(\mathrm{true}\) precisely when an odd number of indices \(i \in s\) satisfy \(f(i) = \mathrm{true}\). Then the sign of this parity equals the product of the individual signs:
Write \(\mathrm{sgn} : \mathrm{Bool} \to \mathbb {R}\) for the Boolean sign map, which sends \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\). Then for all Boolean values \(a, b\),
where \(a \oplus b\) denotes the exclusive-or of \(a\) and \(b\); equivalently, \(\mathrm{sgn}\) is a homomorphism from \((\mathrm{Bool}, \oplus )\) to the multiplicative group \(\{ \pm 1\} \).
Let \(\alpha \) be a finite alphabet, fix a word \(u \colon \{ 1,\dots ,n\} \to \alpha \) of length \(n\), and let \(S \subseteq \{ 1,\dots ,n\} \) be a set of coordinate positions. Then the number of words \(v\) of length \(n\) over \(\alpha \) that differ from \(u\) at exactly the positions in \(S\) (that is, \(v_i \ne u_i\) for \(i \in S\) and \(v_i = u_i\) for \(i \notin S\)) is
Let \(\alpha \) be a finite alphabet, let \(u \colon \mathrm{Fin}\, n \to \alpha \) be a word of length \(n\) over \(\alpha \), and let \(i\) be a position. Then the set of symbols \(a \in \alpha \) with \(a \ne u_i\) has cardinality \(\left\lvert \alpha \right\rvert - 1\).
Let \(\alpha \) be a finite alphabet, and let \(u\) be a word of length \(n\) over \(\alpha \), that is, a function \(u \colon \mathrm{Fin}\, n \to \alpha \). Fix a set \(S \subseteq \mathrm{Fin}\, n\) of coordinate positions together with a further position \(i \in \mathrm{Fin}\, n\), and let \(F\) be the support fiber of \(u\) over \(\{ i\} \cup S\), namely the set of words that differ from \(u\) in exactly the positions belonging to \(\{ i\} \cup S\). Since \(i\) lies in \(\{ i\} \cup S\), every word in \(F\) carries at position \(i\) some symbol distinct from \(u_i\), and grouping the words of \(F\) by that symbol partitions \(F\): it equals the disjoint union, taken over all symbols \(a \in \alpha \) with \(a \ne u_i\), of the sets \(\{ v \in F : v_i = a\} \),
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \) with \(\mu \) absolutely continuous with respect to \(\nu \), and write \(\frac{d\mu }{d\nu }\) for the (real-valued) Radon–Nikodym density. Suppose the log-likelihood ratio \(x \mapsto \log \frac{d\mu }{d\nu }(x)\) is \(\mu \)-integrable. Then
where the inner integral on the left is the \(L^{1}(\nu )\)-distance of the density from the constant \(1\).
The density positive integral is \(\mathrm{densityPositiveIntegral}(\mu ,\nu ) := \int _{\mathrm{densityPositiveSet}} (\mathrm{rnDensity}(\mu ,\nu )(x)-1)\, d\nu \), the excess mass of \(\mu \) over \(\nu \) on the region where \(\mu \) dominates.
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), with \(\mu \) absolutely continuous with respect to \(\nu \), and write \(\frac{d\mu }{d\nu }\) for the Radon–Nikodym density of \(\mu \) with respect to \(\nu \). Then the excess mass of the density above \(1\), integrated over the region where the density is at least \(1\), equals half the \(L^1(\nu )\)-distance of the density from \(1\):
Let \(g : X \times Y \to \alpha \) be a two-argument function and let \(S \subseteq X \times Y\) be a fooling set for \(g\), meaning that every rectangle that is monochromatic for \(g\) meets \(S\) in at most one point. Then the deterministic communication complexity of \(g\) satisfies
where \(\left\lvert S\right\rvert \) is the number of elements of \(S\) and both sides are compared in \(\mathbb {N}_\infty \).
The deterministic communication complexity of \(f : X \to Y \to \alpha \) is defined as
The infimum is taken over all deterministic protocols \(p\) that compute \(f\), and the result lives in \(\mathrm{ENat}\) to accommodate the case where no finite protocol exists.
Let \(X\) and \(Y\) be finite nonempty types, let \(\alpha \) be any type, and let \(f : X \to Y \to \alpha \). Then the deterministic communication complexity of \(f\) satisfies
where \(\left\lvert X\right\rvert \) and \(\left\lvert Y\right\rvert \) denote the cardinalities of \(X\) and \(Y\).
Let \(X\) and \(\alpha \) be finite nonempty types, let \(Y\) be an arbitrary type, and let \(f : X \to Y \to \alpha \) be a two-argument function. Then the deterministic communication complexity of \(f\) satisfies
where \(\left\lvert X\right\rvert \) and \(\left\lvert \alpha \right\rvert \) denote the cardinalities of \(X\) and \(\alpha \).
Let \(X\) be an arbitrary type and let \(Y\) and \(\alpha \) be finite nonempty types. Then for every function \(f : X \to Y \to \alpha \), the deterministic communication complexity of \(f\) satisfies
where \(\left\lvert Y\right\rvert \) and \(\left\lvert \alpha \right\rvert \) denote the cardinalities of \(Y\) and \(\alpha \).
Let \(f : X \to Y \to \alpha \) be a two-argument function and let \(n\) be a natural number. Then the deterministic communication complexity \(D(f)\) satisfies \(D(f) \le n\) if and only if there exists a deterministic communication protocol \(p\) over inputs \(X\) and \(Y\) that computes \(f\) and whose communication complexity is at most \(n\).
Let \(X\), \(Y\), and \(\alpha \) be types, let \(f : X \to Y \to \alpha \) be a two-argument function, and let \(n\) be a natural number. Then the deterministic communication complexity satisfies \(D(f) \le n\) if and only if there exists a finite-message protocol \(p\) over inputs \(X\), \(Y\) and outputs \(\alpha \) whose execution computes \(f\), in the sense that \(p\) run on inputs \(x\) and \(y\) returns \(f(x)(y)\) for all \(x\) and \(y\), and whose communication complexity is at most \(n\).
Let \(f : X \to Y \to \alpha \) and let \(p\) be a finite-message protocol over \(X\), \(Y\), \(\alpha \) whose run function equals \(f\). Then for every \(c \in \mathbb {N} \cup \{ \infty \} \) with \(\texttt{FiniteMessage.Protocol.complexity}\, (p) \le c\), the deterministic communication complexity of \(f\) satisfies \(D(f) \le c\). Each of the three public bounds above is a single application of this lemma to an explicit two-message protocol.
A generalized deterministic two-party communication protocol over input types \(X\), \(Y\) and output type \(\alpha \). At each step a player sends an element of an arbitrary finite nonempty type \(\beta \): the alice constructor takes a function \(f : X \to \beta \) and a continuation \(P : \beta \to \mathrm{Protocol}\, X\, Y\, \alpha \), and dually for bob. The output constructor terminates the protocol with a value in \(\alpha \). This inductive type is equivalent to the binary Deterministic.Protocol up to complexity, where a \(\beta \)-valued message costs \(\lceil \log _2 |\beta | \rceil \) bits.
Given a protocol \(p : \mathrm{Protocol}\, X\, Y\, \alpha \) and a continuation \(q : \alpha \to \mathrm{Protocol}\, X\, Y\, \beta \), bind \(p\) \(q\) sequentially composes them: each output leaf \(a\) of \(p\) is replaced by the sub-protocol \(q\, a\), giving a protocol of type \(\mathrm{Protocol}\, X\, Y\, \beta \).
Let \(p\) be a finite-message protocol with inputs of types \(X\), \(Y\) and outputs in a type \(\alpha \), and let \(q\) assign to each output value \(a \in \alpha \) a finite-message protocol \(q\, a\) with inputs of types \(X\), \(Y\) and outputs in a type \(\beta \). Suppose there is a natural number \(c\) such that every continuation has the same communication complexity, that is, \(\mathrm{complexity}(q\, a) = c\) for all \(a \in \alpha \). Then the sequential composition of \(p\) with \(q\) has communication complexity \(\mathrm{complexity}(p) + c\).
Let \(p\) be a finite-message protocol with input types \(X\), \(Y\) and output type \(\alpha \), and let \(q\) assign to each value \(a \in \alpha \) a finite-message protocol \(q(a)\) with the same input types and output type \(\beta \). For any inputs \(x \in X\) and \(y \in Y\), executing the sequential composition of \(p\) with \(q\) agrees with first running \(p\) on \((x,y)\) to obtain an output \(a\) and then running \(q(a)\) on the same inputs:
Given maps \(f_X : X' \to X\) and \(f_Y : Y' \to Y\), the pullback \(p.\mathrm{comap}\, f_X\, f_Y\) is the finite-message protocol over \(X'\), \(Y'\) obtained by precomposing every message function with \(f_X\) or \(f_Y\) respectively, leaving the protocol tree structure unchanged.
Let \(p\) be a finite-message deterministic communication protocol over input types \(X\) and \(Y\) with output type \(\alpha \), and let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be maps of input types. Then the pullback protocol obtained by precomposing every message function of \(p\) with \(f_X\) or \(f_Y\) (as appropriate) has the same communication complexity as \(p\).
Let \(p\) be a finite-message deterministic communication protocol over input types \(X\), \(Y\) with output type \(\alpha \), and let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be maps into those input types. Then for every pair of inputs \(x' : X'\) and \(y' : Y'\), executing the pullback protocol \(p.\mathrm{comap}\, f_X\, f_Y\) on \((x', y')\) yields the same output as executing \(p\) on the transported inputs \((f_X(x'), f_Y(y'))\); that is,
An auxiliary construction: given \(d\) binary query functions \(\mathrm{query}_i : X \to \mathrm{Bool}\) and a family of binary protocols \(Q : (\mathrm{Fin}\, d \to \mathrm{Bool}) \to \mathrm{Protocol}\, X\, Y\, \alpha \), builds a single binary protocol that reads each query bit from Alice in sequence and then runs \(Q\) on the collected bit pattern.
Fix a natural number \(d\), together with binary query functions \(\mathrm{query}_0,\dots ,\mathrm{query}_{d-1} : X \to \mathrm{Bool}\) and a family of deterministic protocols \(Q(\mathrm{bits})\) indexed by the bit patterns \(\mathrm{bits} : \mathrm{Fin}\, d \to \mathrm{Bool}\). Form the protocol that has Alice read the bits \(\mathrm{query}_0(x),\dots ,\mathrm{query}_{d-1}(x)\) in sequence and then run \(Q\) on the collected pattern. Its communication complexity is
the maximum being taken over all \(2^d\) length-\(d\) bit patterns.
Fix a depth \(d \in \mathbb {N}\), a family of binary query functions \(\mathrm{query}_i : X \to \mathrm{Bool}\) for \(i \in \mathrm{Fin}\, d\), and a family of deterministic protocols \(Q(b)\) indexed by bit patterns \(b : \mathrm{Fin}\, d \to \mathrm{Bool}\). Then for every pair of inputs \(x : X\) and \(y : Y\), the protocol that first has Alice read the \(d\) query bits in sequence and then runs \(Q\) on the collected pattern evaluates on \((x,y)\) to the same value as the single protocol \(Q\bigl(i \mapsto \mathrm{query}_i(x)\bigr)\) evaluated on \((x,y)\).
The worst-case communication cost of a finite-message protocol, measured in bits. An output node costs \(0\); an alice or bob node with message type \(\beta \) costs \(\lceil \log _2 |\beta | \rceil \) plus the maximum complexity over all continuations, i.e. \(\lceil \log _2 |\beta | \rceil + \sup _{b \in \beta } \mathrm{complexity}(P\, b)\).
Let \(X\), \(Y\), and \(\alpha \) be types, let \(\beta \) be a finite nonempty type, let \(f : X \to \beta \) be a function, and let \(Q : \beta \to \mathrm{Protocol}\, X\, Y\, \alpha \) assign to each value \(b \in \beta \) a deterministic two-party communication protocol. Then there exists a deterministic protocol \(R\) over the same input types such that, on every pair of inputs \(x : X\) and \(y : Y\), running \(R\) yields the same output as running the protocol \(Q(f(x))\), and the communication complexity of \(R\) equals
where \(\left\lvert \beta \right\rvert \) denotes the cardinality of \(\beta \).
Let \(\iota \) be a finite, nonempty type, let \(f : \iota \to \mathbb {N}\) be a function, and let \(c \in \mathbb {N}\) be a natural number. Then
Given a function \(g : \alpha \to \beta \) and a protocol \(p : \mathrm{Protocol}\, X\, Y\, \alpha \), CommunicationComplexity.Deterministic.FiniteMessage.Protocol.map \(g\) \(p\) is the protocol that runs \(p\) and then applies \(g\) to the output, yielding a protocol of type \(\mathrm{Protocol}\, X\, Y\, \beta \) with the same communication structure as \(p\).
Let \(X\), \(Y\), \(\alpha \), and \(\beta \) be types, let \(g : \alpha \to \beta \) be any function, and let \(p\) be a finite-message deterministic two-party communication protocol with input types \(X\), \(Y\) and output type \(\alpha \). Applying \(g\) to the output of \(p\) — that is, running \(p\) and post-processing its result through \(g\) — yields a protocol whose communication complexity equals that of \(p\):
Let \(g : \alpha \to \beta \) be a function between output types, let \(p\) be a finite-message protocol with outputs in \(\alpha \), and let \(x \in X\) and \(y \in Y\) be inputs. Then
Embeds a binary protocol into the generalized finite-message framework by treating each Boolean message as an element of \(\beta = \mathrm{Bool}\), giving a CommunicationComplexity.Deterministic.Protocol \(X\, Y\, \alpha \) with the same tree structure.
Let \(p\) be a deterministic binary two-party communication protocol over input types \(X\) and \(Y\) with output type \(\alpha \), and let its embedding into the finite-message framework send each Boolean message as an element of \(\{ \mathtt{false},\mathtt{true}\} \). Then the communication complexity of the embedded finite-message protocol equals the communication complexity of \(p\).
Let \(X\) and \(Y\) be the input types of Alice and Bob and let \(\alpha \) be an output type. For every binary deterministic communication protocol \(p\) over \(X\), \(Y\), and \(\alpha \), there is a finite-message protocol \(P\) over the same input and output types that computes the same output on every pair of inputs — that is, \(\mathrm{run}(P) = \mathrm{run}(p)\) as functions of \(x \in X\) and \(y \in Y\) — and has the same communication complexity, \(\mathrm{complexity}(P) = \mathrm{complexity}(p)\).
Let \(X\) and \(Y\) be the input types of Alice and Bob, let \(\alpha \) be an output type, and let \(p\) be a deterministic binary two-party communication protocol producing values in \(\alpha \). Regard \(p\) as a finite-message protocol by treating each of its Boolean messages as an element of the two-element type. Then for all inputs \(x \in X\) and \(y \in Y\), executing this embedded finite-message protocol on \(x\) and \(y\) yields the same value in \(\alpha \) as executing \(p\) directly on \(x\) and \(y\).
Given a family of protocols \(p_i : \mathrm{Protocol}\, X_i\, Y_i\, \alpha _i\) for \(i \in \mathrm{Fin}\, k\), pi \(p\) is the protocol of type \(\mathrm{Protocol}\, (\Pi _i\, X_i)\, (\Pi _i\, Y_i)\, (\Pi _i\, \alpha _i)\) that runs each \(p_i\) on the \(i\)-th components and collects all outputs into a tuple. It is defined by induction on \(k\), using bind, comap, and map.
Let \(k\) be a natural number, and for each index \(i \in \{ 0, 1, \dots , k-1\} \) let \(p_i\) be a finite-message protocol over input types \(X_i\), \(Y_i\) with output type \(\alpha _i\). The \(k\)-fold product protocol runs each \(p_i\) on the \(i\)-th coordinates of a pair of input tuples in \(\prod _i X_i\) and \(\prod _i Y_i\) and collects the outputs into a tuple in \(\prod _i \alpha _i\). Then the communication complexity of the \(k\)-fold product protocol equals the sum of the communication complexities of the factors:
Let \(k \in \mathbb {N}\), and for each index \(i \in \mathrm{Fin}\, k\) let \(p_i\) be a finite-message protocol over input types \(X_i\), \(Y_i\) with output type \(\alpha _i\). Given input tuples \(x = (x_i)_{i}\) and \(y = (y_i)_{i}\), running the \(k\)-fold product protocol on \(x\) and \(y\) yields the tuple whose \(i\)-th coordinate is the result of running \(p_i\) on the components \(x_i\) and \(y_i\); that is,
Given \(p_1 : \mathrm{Protocol}\, X_1\, Y_1\, \alpha _1\) and \(p_2 : \mathrm{Protocol}\, X_2\, Y_2\, \alpha _2\), their product \(p_1.\texttt{prod}\; p_2 : \mathrm{Protocol}\, (X_1 \times X_2)\, (Y_1 \times Y_2)\, (\alpha _1 \times \alpha _2)\) runs \(p_1\) on the first components and \(p_2\) on the second components, pairing the outputs.
Let \(p_1\) be a finite-message protocol over inputs \(X_1\), \(Y_1\) with outputs in \(\alpha _1\), and let \(p_2\) be a finite-message protocol over inputs \(X_2\), \(Y_2\) with outputs in \(\alpha _2\). Then the communication complexity of their product protocol, which runs \(p_1\) on the first input components and \(p_2\) on the second and pairs the two outputs, equals the sum of their individual complexities:
Let \(p_1\) be a finite-message protocol with input types \(X_1\), \(Y_1\) and output type \(\alpha _1\), and let \(p_2\) be a finite-message protocol with input types \(X_2\), \(Y_2\) and output type \(\alpha _2\). For every pair of inputs \(x = (x_1, x_2) \in X_1 \times X_2\) and \(y = (y_1, y_2) \in Y_1 \times Y_2\), running the product protocol of \(p_1\) and \(p_2\) on \((x, y)\) yields the pair of the two componentwise executions:
Given a finite-message protocol \(p\) and inputs \(x : X\), \(y : Y\), run executes \(p\) recursively: at an alice node it evaluates \(f(x)\) and follows the corresponding continuation, at a bob node it evaluates \(f(y)\), and at an output node it returns the stored value.
Given a deterministic finite-message protocol \(p\), the operation CommunicationComplexity.Deterministic.FiniteMessage.Protocol.toPrivateCoin converts it into a private-coin finite-message protocol over arbitrary coin spaces \(\Omega _X\) and \(\Omega _Y\) by ignoring both coin inputs (via comap with \(\mathrm{Prod.snd}\)).
Let \(p\) be a deterministic finite-message protocol over input types \(X\), \(Y\) and output type \(\alpha \), and let \(\Omega _X\), \(\Omega _Y\) be arbitrary types to serve as private-randomness spaces. Form the private-coin finite-message protocol over \(\Omega _X\), \(\Omega _Y\) obtained from \(p\) by having each player ignore their private coin, so that Alice’s message function \(f : X \to \beta \) is replaced by \((\omega , x) \mapsto f(x)\) and Bob’s by \((\omega , y) \mapsto f(y)\), with the protocol tree otherwise unchanged. Then this private-coin protocol has the same communication complexity as \(p\).
Let \(p\) be a deterministic finite-message protocol over input types \(X\) and \(Y\) with output type \(\alpha \), and fix arbitrary private-coin spaces \(\Omega _X\) and \(\Omega _Y\). Form from \(p\) the private-coin protocol over \(\Omega _X\), \(\Omega _Y\) that runs \(p\) on the underlying inputs while discarding both players’ coins. Then for all inputs \(x \in X\), \(y \in Y\) and all coin draws \(\omega _x \in \Omega _X\), \(\omega _y \in \Omega _Y\), executing this private-coin protocol on \(x\), \(y\) with the coins \(\omega _x\), \(\omega _y\) returns exactly the value that the deterministic execution of \(p\) on \(x\), \(y\) produces, regardless of the coins.
A noncomputable function that converts a finite-message protocol \(p\) into a binary protocol CommunicationComplexity.Deterministic.FiniteMessage.Protocol.toProtocol \(p\) with the same run behavior and the same communication complexity, encoding each \(\beta \)-valued message as \(\lceil \log _2 |\beta | \rceil \) bits.
Let \(p\) be a finite-message deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with outputs in \(\alpha \), in which each message is an element of an arbitrary finite nonempty type. Its conversion into a binary protocol, obtained by encoding every \(\beta \)-valued message as \(\lceil \log _2 \left\lvert \beta \right\rvert \rceil \) bits, has communication complexity equal to that of \(p\).
Let \(X\), \(Y\), and \(\alpha \) be types, and let \(p\) be a finite-message two-party protocol over inputs \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \), in which each message ranges over an arbitrary finite nonempty type and a \(\beta \)-valued message is charged \(\lceil \log _2 |\beta | \rceil \) bits. Then there exists a binary deterministic protocol \(P\) over the same input types \(X\), \(Y\) and output type \(\alpha \) — one in which every message is a single bit — that computes the same function, i.e. \(P\) and \(p\) return the same output on every pair of inputs, and has exactly the same communication complexity as \(p\).
Let \(p\) be a finite-message deterministic two-party protocol over input types \(X\) and \(Y\) with output type \(\alpha \), and let \(q\) be the binary protocol obtained by converting \(p\) (encoding each message drawn from a finite nonempty type \(\beta \) as \(\lceil \log _2 \left\lvert \beta \right\rvert \rceil \) bits). Then \(q\) computes the same function as \(p\): for every pair of inputs \(x \in X\) and \(y \in Y\), running \(q\) on \((x, y)\) returns the same value in \(\alpha \) as running \(p\) on \((x, y)\).
Let \(g : X \times Y \to \alpha \) be a two-argument function, and let \(S \subseteq X \times Y\) be a fooling set for \(g\), meaning that \(S\) meets every \(g\)-monochromatic rectangle in at most one point. If the deterministic communication complexity of \(g\) satisfies \(D(g) \le n\) for some \(n \in \mathbb {N}\), then the cardinality of \(S\) satisfies \(\left\lvert S\right\rvert \le 2^{n}\).
Let \(f : X \to Y \to \alpha \) be a two-argument function, let \(D(f) \in \mathbb {N}_\infty \) denote its deterministic communication complexity, and let \(n\) be a natural number. Then \(n \le D(f)\) if and only if every deterministic protocol \(p\) that computes \(f\) has communication complexity at least \(n\).
Let \(g : X \to Y \to \alpha \) be a two-argument function and let \(n\) be a natural number. If every monochromatic rectangle partition of \(g\) consists of strictly more than \(2^n\) parts, then the deterministic communication complexity of \(g\) satisfies \(n + 1 \le D(g)\).
Let \(g : X \times Y \to \alpha \) be a two-argument function, and let \(n\) be a natural number. If the deterministic communication complexity \(D(g)\) satisfies \(D(g) \le n\), then there is a collection \(\mathcal{P}\) of subsets of \(X \times Y\) that is a monochromatic rectangle partition for \(g\) — every member is a rectangle, every member is monochromatic for \(g\), the members cover \(X \times Y\), and distinct members are disjoint — and whose number of parts satisfies \(\left\lvert \mathcal{P}\right\rvert \le 2^n\).
The ceiling of the base-\(2\) logarithm of \(2\) equals \(1\): in Lean’s natural-number ceiling logarithm, \(\lceil \log _2 2 \rceil = 1\). This is the numeric fact used to specialize the additive \(\lceil \log _2 \left\lvert \alpha \right\rvert \rceil \) overhead to \(1\) for Boolean-valued functions.
Let \(f : X \to Y \to \alpha \) be a function and let \(n \in \mathbb {N}\). The one-way deterministic communication complexity of \(f\) is at most \(n\) if and only if there exists a one-way deterministic protocol \(p\) that computes \(f\) and whose bit cost satisfies \(\mathrm{cost}(p) \le n\).
Let \(f : X \to Y \to \alpha \) be a two-argument function whose output type \(\alpha \) is finite and nonempty, and let \(n \in \mathbb {N}\). If the one-way deterministic communication complexity of \(f\) is at most \(n\), then the (interactive) deterministic communication complexity of \(f\) is at most \(n + \lceil \log _2 \left\lvert \alpha \right\rvert \rceil \).
Let \(X\) and \(Y\) be the input types of Alice and Bob, and let \(f : X \to Y \to \mathrm{Bool}\) be a two-party Boolean-valued function. If the one-way deterministic communication complexity of \(f\) is at most \(n\) for some \(n \in \mathbb {N}\), then the (interactive) deterministic communication complexity of \(f\) is at most \(n + 1\).
Let \(f : X \to Y \to \alpha \) be a function and let \(k \in \mathbb {N}\). Then \(k\) is at most the one-way deterministic communication complexity of \(f\) (as an inequality in \(\mathbb {N}_\infty \)) if and only if every one-way deterministic protocol that computes \(f\) has bit cost at least \(k\).
A one-way deterministic communication protocol for inputs \(x \in X\), \(y \in Y\), and output type \(\alpha \) is a structure consisting of a finite, nonempty message codebook Message, Alice’s encoder \(\mathtt{send} : X \to \mathtt{Message}\) that selects a codeword from the codebook based solely on her input, and Bob’s decoder \(\mathtt{decode} : \mathtt{Message} \times Y \to \alpha \) that produces the final output from the received message and his own input.
Given a deterministic one-way protocol \(p\) computing a function \(f : X \to Y \to \alpha \), and a finite probability distribution \(\mu \) on \(X \times Y\), the distributional error of \(p\) with respect to \(\mu \) and \(f\) is the real number
i.e. the \(\mu \)-measure of the set of inputs on which the protocol’s output disagrees with \(f\).
Given a one-way protocol \(p\) with finite output type \(\alpha \), one obtains an interactive finite-message protocol by having Alice send her one-way message and then having Bob encode the decoded output as an index in \(\{ 0,\ldots ,|\alpha |-1\} \) and transmit it.
Let \(p\) be a one-way deterministic communication protocol with inputs in \(X\), \(Y\) and output type \(\alpha \), and suppose \(\alpha \) is finite and nonempty. The finite-message protocol obtained from \(p\) — Alice sends her one-way message, after which Bob transmits the decoded output as an index in \(\{ 0,\ldots ,\left\lvert \alpha \right\rvert -1\} \) — has communication complexity
that is, the bit cost of \(p\) plus \(\lceil \log _2 \left\lvert \alpha \right\rvert \rceil \), where \(\mathrm{Message}\) is the codebook of \(p\).
Let \(p\) be a one-way deterministic communication protocol over input types \(X\), \(Y\) and a finite, nonempty output type \(\alpha \), and let \(x \in X\) and \(y \in Y\). Then executing the finite-message protocol obtained from \(p\) by the embedding into the finite-message model on the inputs \(x\) and \(y\) yields the same element of \(\alpha \) as executing \(p\) directly on \(x\) and \(y\).
A deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) producing a value of type \(\alpha \) is an inductive type with three constructors: an output node carrying the final value, an alice node in which Alice applies a function \(f : X \to \mathrm{Bool}\) to her input and branches on the resulting bit, and a bob node in which Bob does the same with a function \(g : Y \to \mathrm{Bool}\).
Let \(X\), \(Y\), and \(\alpha \) be types, let \(f : X \to \mathrm{Bool}\), and let \(P\) assign to each bit \(b \in \mathrm{Bool}\) a deterministic communication protocol \(P(b)\) over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \). Consider the protocol whose root is an Alice node: on inputs \(x\) and \(y\) it applies \(f\) to \(x\) and continues with \(P(f(x))\). Then there exists a protocol \(q\) over input types \(Y\) (Alice) and \(X\) (Bob) with output type \(\alpha \) such that, for all \(x \in X\) and \(y \in Y\), executing \(q\) on inputs \(y\) and \(x\) returns the same value of \(\alpha \) as executing the original Alice-rooted protocol on inputs \(x\) and \(y\), and moreover \(q\) has the same communication complexity as that protocol.
Let \(p\) be a protocol on \(X \times Y\), let \(A \subseteq X\), \(B \subseteq Y\), let \(f : X \to \{ 0,1\} \) be the bit Alice sends, and let \(b\) be a bit. If \(R\) is a leaf rectangle of \(p\) run on the restricted input set \((A \cap \{ x \mid f(x) = b\} ) \times B\), then every input \((x, y) \in R\) satisfies \(f(x) = b\).
Let \(p\) be a protocol on \(X \times Y\), let \(A \subseteq X\), \(B \subseteq Y\), let \(f : Y \to \{ 0,1\} \) be the bit Bob sends, and let \(b\) be a bit. If \(R\) is a leaf rectangle of \(p\) run on the restricted input set \(A \times (B \cap \{ y \mid f(y) = b\} )\), then every input \((x, y) \in R\) satisfies \(f(y) = b\). This is the mirror image of CommunicationComplexity.Deterministic.Protocol.aux_alice_bit.
Let \(p\) be a deterministic two-party communication protocol with input types \(X\) (Alice) and \(Y\) (Bob), and let \(A \subseteq X\) and \(B \subseteq Y\). Then the number of distinct rectangles among the auxiliary leaf rectangles of \(p\) relative to the constraint \(A \times B\) is at most \(2^{c}\), where \(c\) is the communication complexity of \(p\).
Let \(S_0, S_1\) be sets in a type \(\beta \) and let \(c_0, c_1 \in \mathbb {N}\). If \(\left\lvert S_0\right\rvert \le 2^{c_0}\) and \(\left\lvert S_1\right\rvert \le 2^{c_1}\), then
where \(\left\lvert \cdot \right\rvert \) denotes the number of elements of a set.
Let \(p\) be a deterministic two-party communication protocol with Alice’s inputs drawn from a set \(X\), Bob’s inputs from a set \(Y\), and outputs in a type \(\alpha \), and let \(A \subseteq X\) and \(B \subseteq Y\). Then the rectangle \(A \times B\) is contained in the union of the auxiliary leaf rectangles of \(p\) relative to the constraint \(A \times B\); that is, every pair \((x, y) \in A \times B\) lies in at least one rectangle of this collection.
Let \(p\) be a deterministic communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) producing values of type \(\alpha \), and let \(A \subseteq X\) and \(B \subseteq Y\). If \(R\) and \(S\) are two distinct rectangles in the collection of auxiliary leaf rectangles of \(p\) relative to \(A\) and \(B\), then \(R\) and \(S\) are disjoint as subsets of \(X \times Y\).
Let \(p, q\) be protocols on \(X \times Y\), let \(A \subseteq X\), \(B \subseteq Y\), let \(f : X \to \{ 0,1\} \), and let \(b \neq b'\) be distinct bits. If \(R\) is a leaf rectangle of \(p\) on \((A \cap \{ x \mid f(x) = b\} ) \times B\) and \(S\) is a leaf rectangle of \(q\) on \((A \cap \{ x \mid f(x) = b'\} ) \times B\), then \(R\) and \(S\) are disjoint.
Let \(p, q\) be protocols on \(X \times Y\), let \(A \subseteq X\), \(B \subseteq Y\), let \(f : Y \to \{ 0,1\} \), and let \(b \neq b'\) be distinct bits. If \(R\) is a leaf rectangle of \(p\) on \(A \times (B \cap \{ y \mid f(y) = b\} )\) and \(S\) is a leaf rectangle of \(q\) on \(A \times (B \cap \{ y \mid f(y) = b'\} )\), then \(R\) and \(S\) are disjoint. This is the mirror image of CommunicationComplexity.Deterministic.Protocol.aux_disjoint_alice_of_ne.
Let \(X\), \(Y\), and \(\alpha \) be types, let \(p\) be a deterministic two-party communication protocol over inputs \(X\) (Alice) and \(Y\) (Bob) with outputs in \(\alpha \), and let \(A \subseteq X\) and \(B \subseteq Y\). Then the collection of leaf rectangles induced by \(p\) when its reachable inputs are constrained to \(A \times B\) is finite.
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob), and let \(A \subseteq X\) and \(B \subseteq Y\). Then every set belonging to the auxiliary leaf-rectangle collection of \(p\) relative to the constraint \(A \times B\) is a rectangle; that is, if \(R \subseteq X \times Y\) is one of these leaf sets, then \(R = A' \times B'\) for some \(A' \subseteq X\) and \(B' \subseteq Y\).
Let \(p\) be a deterministic communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with outputs in \(\alpha \), and let \(A \subseteq X\) and \(B \subseteq Y\). If \(R\) is one of the auxiliary leaf rectangles of \(p\) constrained to \(A \times B\), and \((x, y)\) and \((x', y')\) both belong to \(R\), then running \(p\) yields the same output at both inputs; that is, \(p(x, y) = p(x', y')\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob), and let \(A \subseteq X\) and \(B \subseteq Y\). Then every rectangle \(R\) in the collection of auxiliary leaf rectangles of \(p\) constrained to \(A \times B\) satisfies \(R \subseteq A \times B\).
Let \(p\) be a deterministic two-party communication protocol, and let \(n\) be a natural number with \(n \gt 1\). Write \(\ell (q)\) for the number of leaves of a protocol \(q\), and suppose that \(3\, \ell (p) \ge 2n\). Then there is a subprotocol \(s\) of \(p\) (a rooted subtree of \(p\)) whose leaf count satisfies
Let \(p\) be a protocol, let \(P(\mathrm{false})\) and \(P(\mathrm{true})\) be two protocols that are both subprotocols of \(p\), and let \(n\) be a natural number with \(3\, (\ell (P(\mathrm{false})) + \ell (P(\mathrm{true}))) \ge 2n\), where \(\ell (\cdot )\) denotes the number of leaves. Suppose further that for each bit \(b\) with \(3\, \ell (P(b)) \ge 2n\) there is a subprotocol \(s\) of \(P(b)\) with \(n \le 3\, \ell (s) \lt 2n\). Then there is a subprotocol \(s\) of \(p\) with \(n \le 3\, \ell (s) \lt 2n\).
Let \(p\) be a deterministic two-party communication protocol, and for any such protocol \(q\) write \(\ell (q)\) for its number of leaves (its output nodes). Suppose \(\ell (p) \gt 1\). Then there is a protocol \(s\) that is a rooted subtree of \(p\) whose number of leaves satisfies
so that \(s\) carries at least one-third and strictly fewer than two-thirds of the leaves of \(p\).
Let \(X\) and \(Y\) be the input types of Alice and Bob and let \(\alpha \) be the output type. Fix a function \(f : Y \to \mathrm{Bool}\) together with a family of deterministic communication protocols \(P(\mathtt{false})\) and \(P(\mathtt{true})\), and let \(R\) be the protocol whose root is a bob node: on inputs \(x\) (Alice) and \(y\) (Bob) it computes the bit \(f(y)\) and then runs \(P(f(y))\). Then there is a deterministic communication protocol \(q\) with the roles of Alice and Bob interchanged — Alice’s input drawn from \(Y\) and Bob’s from \(X\) — such that for all \(x\) and \(y\), running \(q\) on Alice-input \(y\) and Bob-input \(x\) yields the same output value as running \(R\) on Alice-input \(x\) and Bob-input \(y\), and moreover \(q\) has the same communication complexity as \(R\).
For every deterministic two-party communication protocol \(p\) with input types \(X\) and \(Y\) and output type \(\alpha \), the number of syntactic transcripts of \(p\) is at most \(2\) raised to its communication complexity:
Given functions \(f_X : X' \to X\) and \(f_Y : Y' \to Y\), comap pulls back a protocol \(p\) over \(X \times Y\) to a protocol over \(X' \times Y'\) by pre-composing every message function with \(f_X\) or \(f_Y\) as appropriate, leaving output nodes unchanged.
Let \(p\) be a deterministic two-party communication protocol with Alice’s inputs drawn from \(X\), Bob’s inputs from \(Y\), and outputs in \(\alpha \), and let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be maps re-indexing the two input types. Then the pull-back protocol over \(X'\) and \(Y'\) obtained by pre-composing each of Alice’s message functions with \(f_X\) and each of Bob’s message functions with \(f_Y\) has the same communication complexity as \(p\).
Let \(p\) be a deterministic two-party communication protocol with Alice’s inputs drawn from \(X\) and Bob’s from \(Y\), producing values in \(\alpha \), and let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be arbitrary maps of input types. Then for all \(x' \in X'\) and \(y' \in Y'\), running the pull-back of \(p\) along \(f_X\) and \(f_Y\) on the inputs \(x'\) and \(y'\) yields the same output as running \(p\) itself on the transported inputs \(f_X(x')\) and \(f_Y(y')\).
The communication complexity of a protocol \(p\) is the worst-case total number of bits exchanged: an output node costs \(0\), while an alice or bob node costs \(1\) plus the maximum complexity of the two sub-protocols reached by the bit \(\mathtt{false}\) and the bit \(\mathtt{true}\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) and \(Y\) with output type \(\alpha \). If \(p\) has exactly one leaf, that is, if the number of output nodes of \(p\) equals \(1\), then the communication complexity of \(p\) is \(0\).
Let \(s\) and \(p\) be deterministic two-party communication protocols, and suppose \(s\) occurs as a rooted subtree of \(p\), witnessed by a subprotocol path recording the branching choices from the root of \(p\) down to the root of \(s\). If \(s\) has strictly fewer leaves than \(p\), then deleting the branch leading to \(s\) yields a protocol: the splicing operation along this path is defined and produces some protocol \(q\).
Let \(s\) and \(p\) be deterministic two-party communication protocols, and suppose that \(s\) occurs as a rooted subprotocol of \(p\), witnessed by a path recording the sequence of left/right branchings from the root of \(p\) down to the root of \(s\). If \(s\) has strictly fewer leaves (output nodes) than \(p\), then splicing out the branch leading to \(s\) does not collapse the whole protocol: the result of the deletion is some protocol rather than the empty result.
Let \(s\) be a subprotocol of \(r\) via a subprotocol path \(h_s\), let \(P\) be a family of children with \(P(b) = r\), and let \(N\) be a node constructor whose leaf count is the sum of its two children’s leaf counts. Let \(Q\) be a rebuilding map with \(Q(q_c)(b) = q_c\) and \(Q(q_c)(\neg b) = P(\neg b)\), and suppose \(q\) is the result of deleting along the path inside branch \(b\): \(q = P(\neg b)\) if the deletion of \(h_s\) returns nothing, and \(q = N(Q(q_c))\) if it returns \(q_c\). Assume (inductively) that whenever the deletion of \(h_s\) returns some \(q'\) one has \(\ell (q') = \ell (r) - \ell (s)\). Then
Let \(p\) be a deterministic communication protocol and let \(s\) be a rooted subtree of \(p\), witnessed by a subprotocol path from the root of \(p\) down to the root of \(s\). Write \(\ell (r)\) for the number of leaves (output nodes) of a protocol \(r\). If deleting this branch splices out a proper subtree — that is, if the deletion produces a protocol \(q\) rather than the empty result of the root case — then
where the subtraction is truncated subtraction on \(\mathbb {N}\).
Let \(s\) be a subprotocol of \(r\) via a subprotocol path \(h_s\), let \((x,y)\) be an input pair, let \(P\) be a family of children with \(P(b) = r\), and let \(N\) be a node constructor such that running \(N(Q)\) on \((x,y)\) is the same as running \(Q(c)\) on \((x,y)\) for a fixed bit \(c\) and every family \(Q\). Let \(Q\) be a rebuilding map with \(Q(q_c)(b) = q_c\) and \(Q(q_c)(\neg b) = P(\neg b)\), and suppose \(q\) is the result of deleting along the path inside branch \(b\): \(q = P(\neg b)\) if the deletion of \(h_s\) returns nothing, and \(q = N(Q(q_c))\) if it returns \(q_c\). Assume that if \((x,y)\) reaches \(s\) along the path then \(c \ne b\), and (inductively) that whenever the deletion of \(h_s\) returns some \(q'\) and \((x,y)\) does not reach \(s\) along the path, \(q'\) runs like \(r\) on \((x,y)\). Then \(q\) and \(N(P)\) produce the same output on \((x,y)\).
Let \(p\) be a deterministic two-party communication protocol on input types \(X\) (Alice) and \(Y\) (Bob), and let \(s\) be a rooted subprotocol of \(p\) witnessed by a subprotocol path from \(s\) to \(p\). Suppose that deleting this path — splicing out the branch of \(p\) leading to \(s\) — yields a protocol \(q\) (rather than being undefined, which happens only when \(s\) is the whole of \(p\)). Then for every pair of inputs \((x,y)\) that does not reach the subprotocol \(s\) along the path — that is, \(x\) is not a reachable Alice input or \(y\) is not a reachable Bob input for that path — executing \(q\) on \((x,y)\) produces the same output as executing \(p\) on \((x,y)\).
Given a deterministic protocol \(p : \mathrm{Protocol}\, X\, Y\, \alpha \), a finite probability space \(\mu \) on \(X \times Y\), and a target function \(f : X \to Y \to \alpha \), the distributional error \(p.\mathrm{distributionalError}(\mu , f) \in \mathbb {R}\) is the \(\mu \)-probability that the output of \(p\) disagrees with \(f\), i.e. \(\mu \bigl(\{ (x,y) \mid p.\mathrm{run}(x,y) \ne f(x,y)\} \bigr)\).
Let \(s\) and \(p\) be deterministic two-party communication protocols over input types \(X\) (Alice) and \(Y\) (Bob) producing values in \(\alpha \), and suppose \(s\) is a rooted subtree of \(p\), as recorded by a subprotocol path from the root of \(p\) down to the root of \(s\). If splicing that path out of \(p\) produces no protocol at all, then \(s\) is in fact all of \(p\); that is, \(s = p\).
Let \(s\) and \(p\) be deterministic two-party communication protocols over the same input types \(X\), \(Y\) and output type \(\alpha \), and suppose that \(s\) is a subprotocol of \(p\), i.e. a rooted subtree of \(p\). Write \(L(q)\) for the number of leaves (output nodes) of a protocol \(q\). Erasing \(s\) from \(p\) — collapsing the subtree \(s\) to a single leaf carrying its canonical (leftmost) output value — yields a protocol whose number of leaves is
where the subtraction is truncated natural-number subtraction.
Let \(p\) be a deterministic two-party communication protocol, and let \(s\) be a subprotocol of \(p\) (a rooted subtree of \(p\)). Form the erased protocol \(\mathrm{erase}(s,p)\), obtained from \(p\) by collapsing the subtree \(s\) to a single leaf. Then for every pair of inputs \(x\) (Alice) and \(y\) (Bob) that does not reach \(s\), running the erased protocol and running \(p\) produce the same output:
Let \(p\) be a deterministic two-party communication protocol, and let \(s\) be a subprotocol of \(p\), witnessed by a subprotocol path recording the sequence of branch choices leading from the root of \(p\) down to the root of \(s\). Form the protocol obtained from \(p\) by collapsing the entire subtree \(s\) to a single output leaf carrying its canonical output value. Writing \(\ell (q)\) for the number of leaves (output nodes) of a protocol \(q\), the collapsed protocol has \(\ell (p) - \ell (s) + 1\) leaves, where the subtraction is truncated subtraction of natural numbers.
Let \(P\) assign a protocol to each bit, let \(b\) be a bit, and let \(s\) be a subprotocol of \(P(b)\) witnessed by a subprotocol path. If \(e = \ell (P(b)) - \ell (s) + 1\) (truncated subtraction on \(\mathbb {N}\)), then replacing the leaf count of branch \(b\) by \(e\) in the sum of the two children’s leaf counts gives
i.e. \(\ell (P(\mathrm{false})) + e = \ell (P(\mathrm{false})) + \ell (P(\mathrm{true})) - \ell (s) + 1\) when \(b = \mathrm{true}\), and \(e + \ell (P(\mathrm{true}))\) equals the same quantity when \(b = \mathrm{false}\).
Let \(p\) be a deterministic two-party communication protocol over Alice’s inputs \(X\) and Bob’s inputs \(Y\), let \(s\) be a rooted subprotocol of \(p\) recorded by a subprotocol path from \(s\) to \(p\), and let \(p'\) be the protocol obtained from \(p\) by collapsing the subtree \(s\) to a single leaf carrying the canonical output of \(s\) (its leftmost leaf value). Suppose an input pair \((x,y)\), with \(x \in X\) and \(y \in Y\), does not reach \(s\): it fails to be consistent with all of Alice’s and Bob’s branch choices along the path from the root of \(p\) down to the root of \(s\). Then running \(p'\) on \((x,y)\) produces the same output as running \(p\) on \((x,y)\).
Let \(s\) be a subprotocol of \(r\) via a subprotocol path \(h_s\), let \((x,y)\) be an input pair, and let \(N\) be a node constructor (an Alice or Bob node) such that for every family \(Q\) of children, running \(N(Q)\) on \((x,y)\) is the same as running \(Q(c)\) on \((x,y)\) for a fixed bit \(c\). Let \(P\) be a family of children with \(P(b) = r\), and let \(Q\) be the family obtained from \(P\) by replacing the branch \(b\) by the erasure of \(h_s\) while keeping the sibling \(Q(\neg b) = P(\neg b)\). Assume that if \((x,y)\) reaches \(s\) along the path then \(c \ne b\), and that (inductively) if \((x,y)\) does not reach \(s\) along the path then the erasure of \(h_s\) runs like \(r\) on \((x,y)\). Then \(N(Q)\) and \(N(P)\) produce the same output on \((x,y)\).
For every deterministic two-party communication protocol \(p\) over input types \(X\) and \(Y\) with outputs in \(\alpha \), there exists a protocol \(q\) that computes the same function — that is, \(q\) and \(p\) return the same value on every pair of inputs \((x,y)\) — and whose communication complexity \(c(q)\) satisfies
where \(\ell (p)\) denotes the number of leaves (output nodes) of \(p\).
The inductive proposition \(\texttt{IsSubprotocol}\; s\; p\) asserts that protocol \(s\) is a rooted subtree of protocol \(p\). It is generated by: reflexivity (\(s\) is a subtree of itself), and three step cases recording that \(s\) is a subtree of one of the two children of an Alice-node or the false-child of a Bob-node, and hence a subtree of the parent.
Let \(s\), \(t\), and \(u\) be deterministic two-party communication protocols over the same input types and output type. If \(s\) is a subprotocol of \(t\) and \(t\) is a subprotocol of \(u\), then \(s\) is a subprotocol of \(u\). In other words, the subprotocol relation is transitive.
The leaf rectangles of a protocol \(p : \texttt{Protocol}\, X\, Y\, \alpha \) form the collection \(\texttt{leafRectangles}(p) = \texttt{leafRectanglesAux}(p, X, Y)\), i.e. the set of input rectangles induced by the leaves of \(p\) over the full input space \(X \times Y\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) and \(Y\) producing values in \(\alpha \), and let \(c\) be its communication complexity. Then the number of leaf rectangles of \(p\) is at most \(2^{c}\).
Let \(p\) be a deterministic communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) producing a value of some type \(\alpha \). Then the leaf rectangles of \(p\) cover the entire input space: the union of all rectangles in the collection of leaf rectangles of \(p\) equals \(X \times Y\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) and \(Y\) producing values in \(\alpha \), and let \(R\) and \(S\) be two of its leaf rectangles, that is, two of the input rectangles induced by the leaves of \(p\) over \(X \times Y\). If \(R \neq S\), then \(R\) and \(S\) are disjoint as subsets of \(X \times Y\).
For any deterministic two-party communication protocol \(p\) over input types \(X\) (Alice) and \(Y\) (Bob) producing values in \(\alpha \), the collection of leaf rectangles of \(p\) — the input rectangles in \(X \times Y\) induced by the leaves of \(p\) over the full input space — is finite.
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) and \(Y\) producing outputs in \(\alpha \), and let \(g : X \to Y \to \alpha \). If \(p\) computes \(g\), then the collection of leaf rectangles of \(p\) is a monochromatic rectangle partition of \(X \times Y\) for \(g\): every leaf rectangle is a rectangle and is monochromatic for \(g\), the leaf rectangles cover all of \(X \times Y\), and any two distinct leaf rectangles are disjoint.
Let \(p\) be a deterministic two-party communication protocol with Alice’s inputs drawn from \(X\) and Bob’s from \(Y\), and let \(R\) belong to the collection of leaf rectangles of \(p\) over the full input space \(X \times Y\). Then \(R\) is a rectangle: there exist subsets \(A \subseteq X\) and \(B \subseteq Y\) with \(R = A \times B\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) and \(Y\) with output type \(\alpha \), and let \(g : X \to Y \to \alpha \) be a function that \(p\) computes, meaning that running \(p\) on any pair of inputs \((x, y)\) returns \(g\, x\, y\). Then every leaf rectangle \(R\) of \(p\) is monochromatic for \(g\): the value \(g\, x\, y\) is the same for all \((x, y) \in R\).
For a protocol \(p\) and sets \(A \subseteq X\), \(B \subseteq Y\), \(\texttt{leafRectanglesAux}(p, A, B)\) is the collection of rectangles induced by the leaves of \(p\) when the reachable inputs are constrained to \(A \times B\). It is defined by structural recursion: an output node yields \(\{ A \times B\} \), while Alice (resp. Bob) nodes split \(A\) (resp. \(B\)) according to the branching function and recurse.
Let \(\mu \) be a finite probability space on \(X \times Y\), let \(g : X \times Y \to \mathrm{Bool}\) be a Boolean function, and let \(\gamma \gt 0\) be a real number such that every combinatorial rectangle \(R \subseteq X \times Y\) satisfies \(\left\lvert \mathrm{disc}_\mu (g, R)\right\rvert \le \gamma \), where \(\mathrm{disc}_\mu (g, R) = \mathbb {E}_{(x,y)\sim \mu }[\mathbf{1}_R(x,y)\, \sigma (g(x,y))]\) and \(\sigma (\mathrm{false}) = 1\), \(\sigma (\mathrm{true}) = -1\). Let \(p\) be a deterministic communication protocol with Boolean outputs, and let \(\varepsilon = \mu (\{ (x,y) : p(x,y) \ne g(x,y)\} )\) be its distributional error in computing \(g\) under \(\mu \). If \(1 - 2\varepsilon \gt 0\), then the communication complexity of \(p\) satisfies
Let \(X \times Y\) carry the structure of a finite probability space, let \(p\) be a deterministic two-party communication protocol on inputs \(X\) (Alice) and \(Y\) (Bob) with Boolean output, and let \(R \subseteq X \times Y\). Then \(R\) belongs to the finite set enumerating the leaf rectangles of \(p\) if and only if \(R\) is a leaf rectangle of \(p\).
Let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be maps, let \(R \subseteq X \times Y\), and let \((x', y') \in X' \times Y'\). Then \((x', y')\) lies in the pullback of \(R\) along the product map \(f_X \times f_Y\) if and only if \((f_X(x'), f_Y(y')) \in R\).
Let \(R \subseteq X \times Y\), and write \(\mathrm{swap}(R) = \{ (y,x) \mid (x,y) \in R\} \subseteq Y \times X\) for the set obtained by exchanging the two coordinates of every element of \(R\). Then for any pair \((y, x) \in Y \times X\), one has \((y, x) \in \mathrm{swap}(R)\) if and only if \((x, y) \in R\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \), and let \((x, y) \in X \times Y\) be an input pair. Executing \(p\) on \((x, y)\) yields a syntactic transcript, and the pair \((x, y)\) itself belongs to the input set of that transcript.
Let \(X \times Y\) carry a finite probability space \(\mu \), let \(g : X \to Y \to \mathrm{Bool}\) be a Boolean function, and let \(\gamma \in \mathbb {R}\). If every rectangle \(R \subseteq X \times Y\) satisfies \(\left\lvert \mathrm{disc}_\mu (g, R)\right\rvert \le \gamma \), then \(\gamma \ge 0\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \). Then the number of leaves of \(p\) — that is, the number of output nodes in its tree shape — is strictly positive, and so is at least one.
Let \(\mu \) be a finite probability measure on \(X \times Y\), let \(g : X \times Y \to \mathrm{Bool}\), and let \(\gamma \in \mathbb {R}\) be such that every rectangle \(R \subseteq X \times Y\) satisfies \(\left\lvert \mathrm{disc}_\mu (g, R)\right\rvert \le \gamma \). Then for every deterministic Boolean-valued communication protocol \(p\), writing \(\mathrm{err}\) for the \(\mu \)-probability that \(p\) disagrees with \(g\) (that is, the measure of \(\{ (x,y) : p(x,y) \ne g(x,y)\} \)) and \(c\) for the communication complexity of \(p\),
Let \(s\) and \(p\) be deterministic two-party communication protocols over the same input types \(X\) and \(Y\) and output type \(\alpha \). If \(s\) is a subprotocol of \(p\) — that is, the subprotocol predicate holds of the pair \((s, p)\) — then the type of subprotocol paths from \(s\) to \(p\) is nonempty: there exists a data-carrying witness recording the sequence of branch choices leading from the root of \(p\) down to the root of \(s\).
Given maps \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) and a set \(R \subseteq X \times Y\), \(\texttt{preimageInputSet}(f_X, f_Y, R)\) is the set \(\{ (x', y') \in X' \times Y' \mid (f_X(x'), f_Y(y')) \in R\} \), the pullback of \(R\) along the product map \(f_X \times f_Y\).
Let \(p\) be a deterministic communication protocol over input types \(X\) and \(Y\) producing values in \(\alpha \), and let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be arbitrary maps. Write \(p'\) for the protocol over \(X' \times Y'\) obtained from \(p\) by pre-composing every message function with \(f_X\) or \(f_Y\) as appropriate. If \(R \subseteq X \times Y\) is a leaf rectangle of \(p\), then its pullback \(\{ (x', y') \in X' \times Y' \mid (f_X(x'), f_Y(y')) \in R\} \) along the product map \(f_X \times f_Y\) is a leaf rectangle of \(p'\).
Let \(p\) be a deterministic communication protocol over input types \(X\) and \(Y\), let \(A \subseteq X\) and \(B \subseteq Y\), and let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be arbitrary maps. If \(R \subseteq X \times Y\) is one of the leaf rectangles of \(p\) relative to the constraint \(A \times B\), then its pullback along \(f_X \times f_Y\), namely \(\{ (x', y') \in X' \times Y' \mid (f_X(x'), f_Y(y')) \in R\} \), is one of the leaf rectangles of the pulled-back protocol \(\mathrm{comap}(p, f_X, f_Y)\) relative to the constraint \(f_X^{-1}(A) \times f_Y^{-1}(B)\).
The noncomputable variant of prunePath taking an IsSubprotocol proof: prune hsp hlt applies CommunicationComplexity.Deterministic.Protocol.prunePath to the classically chosen path for hsp, given the strictness condition \(s.\mathrm{numLeaves} \lt p.\mathrm{numLeaves}\).
Let \(p\) be a deterministic two-party communication protocol and let \(s\) be a subprotocol of \(p\), that is, a rooted subtree of \(p\). Write \(L(q)\) for the number of leaves (output nodes) of a protocol \(q\), and suppose that \(L(s) \lt L(p)\). Then the protocol obtained by pruning \(s\) out of \(p\) has exactly \(L(p) - L(s)\) leaves.
Let \(s\) and \(p\) be deterministic two-party communication protocols over Alice’s inputs \(X\), Bob’s inputs \(Y\), and output type \(\alpha \), and suppose \(s\) is a subprotocol of \(p\) (a rooted subtree) whose leaf count is strictly smaller than that of \(p\). Form the pruned protocol obtained by deleting the subtree \(s\) from \(p\). Then for every pair of inputs \(x \in X\) and \(y \in Y\) that does not reach \(s\) — that is, \((x,y)\) fails to be consistent with all of the branching choices along the path from the root of \(p\) down to \(s\) — the pruned protocol and \(p\) produce the same output on \((x,y)\): running the pruned protocol on \(x\) and \(y\) yields exactly \(p\)’s output on \(x\) and \(y\).
Let \(s\) and \(p\) be deterministic two-party communication protocols, and let \(s\) lie along a subprotocol path inside \(p\), so that \(s\) occurs as a rooted subtree of \(p\). If \(s\) has strictly fewer leaves than \(p\), that is, if the number of leaves of \(s\) is less than the number of leaves of \(p\), then the protocol obtained by pruning this path from \(p\) has exactly \(\text{numLeaves}(p) - \text{numLeaves}(s)\) leaves.
Let \(p\) be a deterministic two-party communication protocol with Alice inputs in \(X\) and Bob inputs in \(Y\), and let \(s\) be a rooted subprotocol of \(p\), witnessed by a subprotocol path from the root of \(p\) down to the root of \(s\). Suppose \(s\) has strictly fewer leaves than \(p\), so that pruning this path yields a protocol \(p'\). Then for every input pair \((x,y)\) that does not reach \(s\) along this path — that is, \(x\) is not among the Alice inputs consistent with all of Alice’s branching choices on the path, or \(y\) is not among the Bob inputs consistent with all of Bob’s choices on it — the pruned protocol \(p'\) and the original protocol \(p\) return the same value on \((x,y)\).
Let \(p\) be a deterministic communication protocol and let \(s\) be a rooted subprotocol of \(p\), recorded by a subprotocol path running from the root of \(p\) down to the root of \(s\). Suppose that \(s\) has strictly fewer leaves than \(p\). Then splicing this path out of \(p\) does not fall into the empty root case, in which the subprotocol coincides with the whole protocol, but instead returns a genuine protocol; and that protocol is exactly the pruned protocol determined by the path together with this leaf-count hypothesis.
Let \(X\), \(Y\), and \(\alpha \) be types, let \(p\) be a deterministic two-party communication protocol over Alice’s inputs \(X\) and Bob’s inputs \(Y\) with outputs in \(\alpha \), and suppose \(s\) is a rooted subtree of \(p\). Then the set of input pairs \((x, y) \in X \times Y\) that reach \(s\) — those for which \(x\) lies in Alice’s reach set and \(y\) lies in Bob’s reach set — is a rectangle: it factors as a product \(A \times B\) for some \(A \subseteq X\) and \(B \subseteq Y\).
\(\texttt{reachesPath}\; \texttt{hsp}\; x\; y\) is the proposition that the input pair \((x, y)\) causes protocol \(p\) to reach the subprotocol \(s\) indicated by the path hsp; that is, \(x \in \texttt{reachXPath}\; \texttt{hsp}\) and \(y \in \texttt{reachYPath}\; \texttt{hsp}\).
Let \(s\) and \(p\) be deterministic two-party communication protocols over Alice’s input type \(X\) and Bob’s input type \(Y\), and let \(\pi \) be a subprotocol path witnessing that \(s\) is a rooted subtree of \(p\). Then the set of input pairs
is a rectangle. Explicitly, it factors as \(A \times B\), where \(A \subseteq X\) is the set of Alice inputs consistent with every one of Alice’s branching choices along \(\pi \) and \(B \subseteq Y\) is the set of Bob inputs consistent with every one of Bob’s branching choices along \(\pi \).
Let \(p\) be a deterministic two-party communication protocol with Alice inputs in \(X\) and Bob inputs in \(Y\), and let \(\pi \) be a subprotocol path witnessing that a subprotocol \(s\) is a rooted subtree of \(p\). If deleting the path \(\pi \) from \(p\) returns no protocol — which occurs precisely when the path is trivial, so that \(s = p\) — then every input pair \((x, y) \in X \times Y\) reaches \(s\) along \(\pi \), meaning \(x\) is consistent with all of Alice’s choices along \(\pi \) and \(y\) is consistent with all of Bob’s choices along \(\pi \).
Let \(p\) be a deterministic two-party communication protocol on input types \(X\) and \(Y\) with output type \(\alpha \), and let \(g : X \to Y \to \alpha \) be a function that \(p\) computes. Then the leaf rectangles of \(p\) form a monochromatic rectangle partition of \(X \times Y\) for \(g\); that is, they are pairwise disjoint rectangles covering \(X \times Y\), each of which is monochromatic for \(g\). Moreover, the number of leaf rectangles is at most \(2^{c}\), where \(c\) is the communication complexity of \(p\).
Let \(p\) be a deterministic Boolean two-party communication protocol on input types \(X\) (Alice) and \(Y\) (Bob), and let \(R \subseteq X \times Y\) be any set. Define the rectangle sign of \(p\) on \(R\) to be \(+1\) when \(p\) outputs \(\mathrm{false}\) at every point of \(R\) and \(-1\) otherwise. Then the rectangle sign has absolute value \(1\).
Let \(p\) be a deterministic two-party communication protocol with Boolean output, over input types \(X\) (Alice) and \(Y\) (Bob). If \(R\) is a leaf rectangle of \(p\) and \((x, y)\) is any point of \(R\), then the rectangle sign of \(R\) agrees with the Boolean sign of the value that \(p\) computes on \((x, y)\): it equals \(+1\) when \(p\) outputs \(\mathrm{false}\) at \((x, y)\) and \(-1\) when \(p\) outputs \(\mathrm{true}\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with outputs in \(\alpha \), and let \((x, y)\) and \((x', y')\) be two input pairs in \(X \times Y\). If executing \(p\) on \((x, y)\) produces the same syntactic transcript as executing \(p\) on \((x', y')\), then \(p\) returns the same output on both inputs: the value obtained by running \(p\) on \((x, y)\) equals the value obtained by running \(p\) on \((x', y')\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \), and let \((x, y) \in X \times Y\) be an input pair. Running \(p\) on \(x\) and \(y\) produces the same value as first executing \(p\) on \((x,y)\) to obtain its syntactic transcript and then reading off the output value attached to the leaf of that transcript.
Given a deterministic communication protocol \(p\) over input types \(X\) and \(Y\) with output type \(\alpha \), shape maps \(p\) to its underlying binary tree of type Tree Unit. An output node becomes a leaf (Tree.nil), while an Alice or Bob branching node becomes an internal node (Tree.node) whose two children are the shapes of the two sub-protocols reached by the bit \(\mathtt{false}\) and \(\mathtt{true}\), respectively.
Let \(\beta , \gamma : \{ \mathrm{false}, \mathrm{true}\} \to \mathrm{Type}\) be two Boolean-indexed families of types and let \(g_b : \beta _b \to \gamma _b\) be a map for each \(b \in \{ \mathrm{false}, \mathrm{true}\} \). If every \(g_b\) is injective, then the map \(\Sigma _b\, \beta _b \to \Sigma _b\, \gamma _b\) sending \((b, t)\) to \((b, g_b(t))\) — which keeps the Boolean index and applies the family fibrewise — is injective.
Let \(X\) and \(Y\) be the input types of Alice and Bob, and let \(\mu \) be a finite probability measure on \(X \times Y\). Write \(\sigma \colon \mathrm{Bool} \to \{ \pm 1\} \) for the map sending \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\). For a deterministic communication protocol \(p\) with Boolean output and a target function \(g \colon X \to Y \to \mathrm{Bool}\), the expected signed correlation between the protocol’s output and the target satisfies
where \(p(x,y)\) denotes the value computed by running \(p\) on \((x,y)\) and \(\varepsilon _\mu (p,g) = \mu \bigl(\{ (x,y) \mid p(x,y) \ne g(x,y)\} \bigr)\) is the distributional error of \(p\) against \(g\) under \(\mu \).
Let \(\mu \) be a finite probability space on \(X \times Y\), let \(p\) be a deterministic Boolean communication protocol with input types \(X\) and \(Y\), and let \(g : X \to Y \to \mathrm{Bool}\). Writing \(\sigma \) for the sign map sending \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\), the signed bias of \(p\) against \(g\) satisfies
where the sum runs over the (finitely many) leaf rectangles \(R\) of \(p\), the value \(p(x,y)\) is the output of \(p\) on inputs \(x\) and \(y\), the rectangle sign \(s(p,R)\) equals \(+1\) if \(p\) outputs \(\mathrm{false}\) at every point of \(R\) and \(-1\) otherwise, and \(\mathrm{disc}_\mu (g,R) = \mathbb {E}_{(x,y)\sim \mu }[\mathbf{1}_R(x,y)\cdot \sigma (g(x,y))]\) is the discrepancy of \(g\) on \(R\).
Let \(s, p, q_{\mathrm{in}}, q_{\mathrm{out}}\) be deterministic two-party protocols over input types \(X\) and \(Y\) with outputs in \(\alpha \), where \(s\) is a subprotocol of \(p\) with strictly fewer leaves than \(p\). Suppose \(q_{\mathrm{in}}\) computes the same function as \(s\), and \(q_{\mathrm{out}}\) computes the same function as the protocol obtained from \(p\) by pruning \(s\) away. Then the glued protocol CommunicationComplexity.Deterministic.Protocol.testSubprotocol, which first tests whether the inputs reach \(s\) and then runs \(q_{\mathrm{in}}\) or \(q_{\mathrm{out}}\) accordingly, computes the same function as \(p\) on every pair of inputs \((x,y)\).
Let \(c_{\mathrm{in}}, c_{\mathrm{out}}, m, n \in \mathbb {N}\) satisfy the two inductive bounds \(3^{c_{\mathrm{in}}} \le 2^{c_{\mathrm{in}}} m^2\) and \(3^{c_{\mathrm{out}}} \le 2^{c_{\mathrm{out}}} (n - m)^2\), together with the balance conditions \(3m \le 2n\) and \(3(n - m) \le 2n\). Then, writing \(c = 2 + \max (c_{\mathrm{in}}, c_{\mathrm{out}})\),
Let \(p\) and \(s\) be deterministic two-party communication protocols with Alice’s inputs drawn from \(X\), Bob’s inputs from \(Y\), and outputs in \(\alpha \), and suppose \(s\) is a subprotocol of \(p\), that is, a rooted subtree of \(p\). If an input pair \((x, y)\) reaches \(s\) — meaning \(x\) is consistent with every one of Alice’s branch choices and \(y\) with every one of Bob’s branch choices along the path from the root of \(p\) down to the root of \(s\) — then evaluating \(p\) on \((x, y)\) returns the same output as evaluating \(s\) on \((x, y)\).
Let \(p\) and \(s\) be deterministic two-party communication protocols over Alice’s input type \(X\) and Bob’s input type \(Y\) with output type \(\alpha \), and let \(\pi \) be a subprotocol path witnessing that \(s\) sits at a rooted subtree of \(p\). Suppose the input pair \((x, y)\) reaches \(s\) along \(\pi \), meaning \(x\) lies in the set of Alice inputs consistent with all of Alice’s branching choices on \(\pi \) and \(y\) lies in the set of Bob inputs consistent with all of Bob’s branching choices on \(\pi \). Then executing \(p\) and executing \(s\) on \((x, y)\) yield the same output.
The inductive type \(\texttt{SubprotocolPath}\; s\; p\) is a data-carrying witness that \(s\) is a rooted subtree of \(p\), recording the sequence of left/right choices made from the root of \(p\) down to the root of \(s\). It mirrors the constructors of IsSubprotocol but lives in Type rather than Prop, enabling structural recursion over the path.
Let \(s\) and \(p\) be deterministic two-party communication protocols over the same input types, and suppose there is a subprotocol path from \(s\) to \(p\), witnessing that \(s\) occurs as a rooted subtree of \(p\). Then \(s\) has at most as many leaves as \(p\); that is, the number of leaves of \(s\) is less than or equal to the number of leaves of \(p\).
Let \(X\) and \(Y\) be input types whose product \(X \times Y\) carries a finite probability space, and let \(p\) be a deterministic Boolean two-party protocol over \(X\) (Alice) and \(Y\) (Bob). Write \(\{ R\} \) for the finite collection of leaf rectangles of \(p\), and for each such rectangle assign the sign \(\varepsilon (R) = +1\) if \(p\) outputs \(\mathrm{false}\) at every point of \(R\) and \(\varepsilon (R) = -1\) otherwise. Then for every point \((x,y) \in X \times Y\),
where \(p(x,y)\) denotes the Boolean value returned by running \(p\) on inputs \(x\) and \(y\).
Let \(p\) be a deterministic two-party communication protocol with Alice’s inputs drawn from a type \(X\), Bob’s from a type \(Y\), and outputs in a type \(\alpha \). Let \(p^{\mathrm{swap}}\) be the protocol obtained by interchanging the roles of Alice and Bob throughout — turning every Alice node into a Bob node and vice versa while leaving output nodes unchanged. Then \(p^{\mathrm{swap}}\) has the same communication complexity as \(p\): the worst-case number of bits exchanged is unaffected by the swap.
Let \(p\) be a deterministic two-party communication protocol with Alice’s inputs drawn from \(X\), Bob’s inputs drawn from \(Y\), and outputs in \(\alpha \), and let \(x : X\) and \(y : Y\) be inputs. Then executing the role-swapped protocol on the input pair \((y, x)\) returns the same value as executing \(p\) on \((x, y)\).
Let \(p\) be a deterministic two-party communication protocol with Alice’s inputs drawn from a type \(X\), Bob’s from a type \(Y\), and outputs in a type \(\alpha \). Applying the role swap twice — first exchanging the roles of Alice and Bob throughout \(p\), then exchanging them again — returns the original protocol \(p\) unchanged.
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob), and let \(p^{\mathrm{swap}}\) be the protocol over \(Y \times X\) obtained by exchanging the roles of Alice and Bob throughout. If \(R \subseteq X \times Y\) is one of the leaf rectangles of \(p\), then its coordinate swap \(\{ (y,x) \mid (x,y) \in R\} \subseteq Y \times X\) is one of the leaf rectangles of \(p^{\mathrm{swap}}\).
Let \(p\) be a deterministic communication protocol over input types \(X\) (Alice) and \(Y\) (Bob), and let \(A \subseteq X\) and \(B \subseteq Y\). Denote by \(p^{\mathrm{sw}}\) the role-swapped protocol over \(Y \times X\) obtained by exchanging Alice and Bob throughout, and for a set \(R \subseteq X \times Y\) write \(R^{\mathrm{sw}} = \{ (y,x) \mid (x,y) \in R\} \subseteq Y \times X\) for its coordinate swap. Then for every rectangle \(R\) in the collection of auxiliary leaf rectangles of \(p\) relative to \(A \times B\), the swapped set \(R^{\mathrm{sw}}\) belongs to the collection of auxiliary leaf rectangles of \(p^{\mathrm{sw}}\) relative to \(B \times A\).
Given that \(s\) is a subprotocol of \(p\) and two protocols qIn and qOut, testSubprotocol hsp qIn qOut is a noncomputable protocol that uses two bits of communication to determine whether the input reaches \(s\), then routes to qIn if it does and to qOut otherwise.
Let \(X\) and \(Y\) be the input types of Alice and Bob in a deterministic two-party communication protocol producing values of type \(\alpha \), and suppose \(s\) is a subprotocol of such a protocol \(p\). For any two further protocols \(q_{\mathrm{in}}\) and \(q_{\mathrm{out}}\), the test subprotocol — which spends one bit for Alice and one bit for Bob to decide whether the input reaches \(s\), then continues as \(q_{\mathrm{in}}\) if it does and as \(q_{\mathrm{out}}\) otherwise — has communication complexity
Let \(s\) be a subprotocol of a deterministic communication protocol \(p\), and let \(q_{\mathrm{in}}\) and \(q_{\mathrm{out}}\) be two further protocols on the same input types. If the input pair \((x,y)\) reaches \(s\) — that is, \(x\) lies in Alice’s reach set and \(y\) in Bob’s reach set for \(s\) — then executing the test subprotocol built from \(s\), \(q_{\mathrm{in}}\), and \(q_{\mathrm{out}}\) on \((x,y)\) yields the same output as executing \(q_{\mathrm{in}}\) on \((x,y)\).
Let \(p\) be a deterministic two-party communication protocol over Alice’s input type \(X\) and Bob’s input type \(Y\) with outputs in \(\alpha \), let \(s\) be a subprotocol of \(p\), and let \(q_{\mathrm{in}}\) and \(q_{\mathrm{out}}\) be two further such protocols. Suppose an input pair \((x, y) \in X \times Y\) does not reach \(s\); that is, \(x\) fails to lie in Alice’s reach set for \(s\), or \(y\) fails to lie in Bob’s reach set for \(s\). Then executing the test subprotocol built from \(s\), \(q_{\mathrm{in}}\), and \(q_{\mathrm{out}}\) on \((x, y)\) yields the same output as executing \(q_{\mathrm{out}}\) on \((x, y)\).
For a deterministic protocol \(p\) over input types \(X\) and \(Y\) with output type \(\alpha \), the type \(\mathrm{Transcript}(p)\) records the full message sequence of an execution path: a terminal protocol has a single transcript (unit), while a communication node contributes one Boolean bit together with the transcript of the chosen child protocol.
For a syntactic transcript \(t\) of \(p\), the input set \(\mathrm{inputSet}(t) \subseteq X \times Y\) is the set of all input pairs \((x, y)\) whose execution of \(p\) follows exactly the path described by \(t\). It is defined recursively: for a terminal protocol it is \(X \times Y\), and for a communication node it is the subset of pairs where the current player’s Boolean function matches the bit recorded in \(t\) and the remaining pair belongs to the child transcript’s input set.
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \), and let \(t\) be a syntactic transcript of \(p\). Then the input set of \(t\) — the set of all input pairs \((x, y) \in X \times Y\) whose execution of \(p\) follows exactly the path recorded by \(t\) — is a rectangle: there exist \(A \subseteq X\) and \(B \subseteq Y\) such that this set equals \(A \times B\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \), let \(t\) be a syntactic transcript of \(p\), and let \((x, y) \in X \times Y\) be an input pair. If \((x, y)\) belongs to the input set of \(t\), then executing \(p\) on \((x, y)\) produces exactly the transcript \(t\); that is, \(p.\mathrm{transcript}(x, y) = t\).
Given maps \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) and a syntactic transcript \(t\) of \(p\), transcriptComap \(p\) \(f_X\) \(f_Y\) \(t\) is the corresponding transcript of the pulled-back protocol \(p.\mathrm{comap}(f_X, f_Y)\), obtained by preserving the Boolean message sequence while re-routing through the input maps.
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \), and let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be maps on the input spaces. Consider the pulled-back protocol over \(X' \times Y'\) obtained by pre-composing every message function of \(p\) with \(f_X\) or \(f_Y\) as appropriate. Then the map that sends each syntactic transcript of \(p\) to the corresponding transcript of this pulled-back protocol — preserving the sequence of Boolean messages while re-routing through the input maps — is injective.
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \), and let \(f_X : X' \to X\) and \(f_Y : Y' \to Y\) be functions; write \(p'\) for the pulled-back protocol over \(X' \times Y'\) obtained by pre-composing each of Alice’s and Bob’s message functions with \(f_X\) and \(f_Y\) respectively. Then for every \(x' \in X'\) and \(y' \in Y'\), running \(p\) on the pair \((f_X(x'), f_Y(y'))\) to obtain its syntactic transcript and then re-routing that transcript through \(f_X\) and \(f_Y\) into a transcript of \(p'\) yields exactly the syntactic transcript produced by running \(p'\) directly on \((x', y')\).
Given a syntactic transcript \(t\) of \(p\), transcriptSwap \(t\) is the corresponding transcript of the swapped protocol \(p.\mathrm{swap}\) (where Alice and Bob exchange roles), obtained by preserving the Boolean message sequence while reinterpreting each node as the opposite player.
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob) with output type \(\alpha \), and let \(p'\) be the role-swapped protocol obtained by exchanging the roles of Alice and Bob throughout \(p\). The transcript-swap map, which sends each syntactic transcript of \(p\) to its corresponding transcript of \(p'\) by preserving the Boolean message sequence while reinterpreting each communication node as the opposite player, is injective: distinct transcripts of \(p\) are sent to distinct transcripts of \(p'\).
Let \(p\) be a deterministic two-party communication protocol over input types \(X\) (Alice) and \(Y\) (Bob), and let \(x \in X\) and \(y \in Y\) be inputs. Running \(p\) on the pair \((x,y)\) produces a syntactic transcript, and applying the transcript-swap operation to it — reinterpreting each communication node as the opposite player while preserving the Boolean message sequence — yields exactly the transcript obtained by running the role-swapped protocol (in which Alice and Bob exchange roles) on the reversed pair \((y,x)\).
Let \(X\) be a finite set, \(Y\) a finite set, and \(f : X \to Y \to \mathrm{Bool}\) a Boolean function. If \(\mathcal{P}\) is a monochromatic rectangle partition of \(X \times Y\) for \(f\), then the rank of \(f\) is at most the number of members of \(\mathcal{P}\):
Let \(X\) and \(Y\) be finite types and let \(f : X \to Y \to \mathrm{Bool}\) be a two-party Boolean function, whose rank \(\mathrm{rank}(f)\) is the \(\mathbb {R}\)-rank of the \(0/1\) matrix \(M_f \in \mathbb {R}^{X \times Y}\) with \((M_f)_{x,y} = 1\) exactly when \(f\, x\, y = \mathrm{true}\). If the deterministic communication complexity \(D(f)\) is at most \(n\) for some \(n \in \mathbb {N}\), then \(\mathrm{rank}(f) \le 2^n\).
Let \(X\) be a finite type and \(Y\) a finite type, and let \(f : X \to Y \to \mathrm{Bool}\) be a Boolean function with associated \(0/1\) matrix \(M_f \in \mathbb {R}^{X \times Y}\), whose rank \(\mathrm{rank}(f) = \mathrm{rank}_{\mathbb {R}}(M_f)\) is the rank of \(f\). Then the deterministic communication complexity \(D(f) \in \mathbb {N}_\infty \) of \(f\) satisfies
Let \(X\) be any set and \(Y\) a finite set, and let \(R \subseteq X \times Y\) be a rectangle, meaning \(R = A \times B\) for some \(A \subseteq X\) and \(B \subseteq Y\). Let \(M_R \in \mathbb {R}^{X \times Y}\) be its indicator matrix, with \((M_R)_{x,y} = 1\) when \((x,y) \in R\) and \((M_R)_{x,y} = 0\) otherwise. Then \(\mathrm{rank}(M_R) \le 1\).
Let \(\mu \) be a finite probability space on \(X \times Y\), let \(g : X \to Y \to \mathrm{Bool}\), and let \(S \subseteq X \times Y\). The discrepancy of \(g\) on \(S\) with respect to \(\mu \) is
where \(\sigma : \mathrm{Bool} \to \{ -1,1\} \) is the sign map CommunicationComplexity.boolSign.
Let \(\mu \) be a finite probability space on the product \(X \times Y\), let \(g : X \times Y \to \mathrm{Bool}\), and let \(S \subseteq X \times Y\). Then the discrepancy of \(g\) on \(S\) satisfies
that is, the probability mass inside \(S\) on which \(g\) is \(\mathrm{false}\) minus the probability mass inside \(S\) on which \(g\) is \(\mathrm{true}\).
Let \(X\) and \(Y\) be arbitrary types, let \(g : X \times Y \to \mathrm{Bool}\) (written \(g(x,y)\)), let \(S \subseteq X \times Y\), and let \(\sigma : \mathrm{Bool} \to \mathbb {R}\) send \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\). Then for every point \((x,y) \in X \times Y\),
where each indicator takes the value \(1\) on its set and \(0\) elsewhere.
Let \(x\) be an \(n\)-bit Boolean input, that is, a function assigning a Boolean value to each of the \(n\) coordinates. If \(x\) is not the all-zero input — the constant function taking the value \(\mathrm{false}\) everywhere — then there exists a coordinate \(i\) at which \(x\) takes the value \(\mathrm{true}\).
Let \(\mu \) be a measure on a measurable space \(\Omega \), let \(X\colon \Omega \to \mathbb {R}\), and let \(t \in \mathbb {R}\). Writing \(\Lambda (t)\) for the cumulant generating function of \(X\) under \(\mu \) at \(t\), the functions \(x \mapsto e^{\, t X(x) - \Lambda (t)}\) and \(x \mapsto e^{-\Lambda (t)} \cdot e^{\, t X(x)}\) are equal, i.e. the tilted exponential factors as the constant \(e^{-\Lambda (t)}\) times \(e^{\, tX}\).
A typeclass for a measurable space \(\Omega \) that is simultaneously finite (carries a Fintype instance) and discrete (every subset is measurable). It deliberately does not bundle a measure, so that a single type can be equipped with many different measures.
Let \(\Omega \) be a finite measurable space, and let \(\mu \) and \(\nu \) be two measures on \(\Omega \). Then \(\mu \) is absolutely continuous with respect to \(\nu \), written \(\mu \ll \nu \), if and only if for every point \(\omega \in \Omega \) with \(\nu (\{ \omega \} ) = 0\) one also has \(\mu (\{ \omega \} ) = 0\).
Let \(\Omega \) be a finite measurable space, \(\alpha \) a finite type on which every subset is measurable, \(\mu \) a finite measure on \(\Omega \), and \(Z : \Omega \to \alpha \) and \(f : \alpha \to \mathbb {R}\) any functions. Then
where \(\mu (Z^{-1}(\{ z\} ))\) denotes the real number giving the measure of the fiber \(Z^{-1}(\{ z\} ) = \{ \omega \in \Omega : Z(\omega ) = z\} \).
Let \(\Omega \) be a finite discrete measurable space, and let \(\mu \) and \(\nu \) be finite measures on \(\Omega \). If there is a point \(\omega \in \Omega \) with \(\nu (\{ \omega \} ) = 0\) but \(\mu (\{ \omega \} ) \neq 0\), then the Kullback–Leibler divergence is infinite, \(D_{\mathrm{KL}}(\mu \, \| \, \nu ) = \infty \). Otherwise,
where \(\log \frac{d\mu }{d\nu }\) denotes the log-likelihood ratio of \(\mu \) with respect to \(\nu \), the measures are read as their real values, and \((x)_+\) takes the enclosed real number as a nonnegative extended real, sending any negative value to \(0\).
Let \(\Omega \) be a finite discrete measurable space, and let \(\mu \) and \(\nu \) be finite measures on \(\Omega \). If there is a point \(\omega \in \Omega \) with \(\nu (\{ \omega \} ) = 0\) but \(\mu (\{ \omega \} ) \neq 0\), then the Kullback–Leibler divergence is infinite:
Otherwise,
where each \(\mu (\{ \omega \} )\), \(\nu (\{ \omega \} )\), \(\mu (\Omega )\), \(\nu (\Omega )\) denotes the (real-valued) mass of the indicated set and \((x)_{+} = \max (x, 0)\) is the nonnegative part.
Let \(\Omega \) be a finite discrete measurable space, and let \(\mu \) and \(\nu \) be probability measures on \(\Omega \). If \(\nu \) assigns positive mass to every point, meaning \(\nu (\{ \omega \} ) \neq 0\) for all \(\omega \in \Omega \), then the Kullback–Leibler divergence \(D_{\mathrm{KL}}(\mu \, \| \, \nu )\) is finite.
Let \(\Omega \) be a finite measurable space equipped with a finite measure \(\mu \), let \(\alpha \) be a finite set, let \(Z : \Omega \to \alpha \) be a measurable (\(\alpha \)-valued) map, and let \(S \subseteq \Omega \). Writing \(\mu ^{\mathbb {R}}\) for the real-valued measure and \(\mu [\, \cdot \mid A\, ]\) for the measure conditioned on an event \(A\), we have
so the total mass of \(S\) is recovered by summing, over the values \(z\) of \(Z\), the mass of the fiber \(Z^{-1}(\{ z\} )\) times the conditional mass of \(S\) given that fiber.
Let \(\Omega \) be a finite measurable space — finite and discrete, so that every subset is measurable — and let \(\mu \) be a finite measure on \(\Omega \), writing \(\mu ^{\mathbb {R}}\) for its real-valued mass function. Then for every subset \(S \subseteq \Omega \),
That is, the measure of \(S\) is the sum of the masses of its individual points.
Let \(\Omega \) be a finite discrete measurable space, \(\alpha \) a finite type, \(\mu \) a finite measure on \(\Omega \), \(Z : \Omega \to \alpha \) a function, and \(P\) a predicate on \(\alpha \). Writing \(\mu ^{\mathbb {R}}\) for the real-valued measure obtained from \(\mu \), the measure of the set of points whose image under \(Z\) satisfies \(P\) decomposes as a sum over the fibers of \(Z\):
Let \(\Omega \) be a finite discrete measurable space, and let \(p\) and \(q\) be probability mass functions on \(\Omega \). Then the Kullback–Leibler divergence between the probability measures they induce is
where in the second case the (nonnegative) real sum is read as an element of \([0,\infty ]\).
Let \(\Omega \) be a finite discrete measurable space, and let \(p\) and \(q\) be probability mass functions on \(\Omega \), each regarded as the probability measure it induces. Then the Kullback–Leibler divergence \(D_{\mathrm{KL}}(p \, \| \, q)\), valued in \([0,\infty ]\), equals \(+\infty \) whenever there exists some \(\omega \in \Omega \) with \(q(\omega ) = 0\) and \(p(\omega ) \neq 0\), and otherwise equals the value in \([0,\infty ]\) obtained from
where each mass \(p(\omega )\), \(q(\omega )\) is taken as its real value and negative sums are read as \(0\).
Let \(\Omega \) be a finite measurable space, let \(\mu \) be a probability measure on \(\Omega \), and let \(f : \Omega \to \mathbb {R}\) be any real-valued function. Then
A typeclass bundling a MeasureSpace structure on \(\Omega \) together with a CommunicationComplexity.FiniteMeasureSpace witness and an IsProbabilityMeasure instance for the canonical volume measure. This makes \(\Omega \) simultaneously a finite discrete measurable space and a probability space.
Let \(\Omega \) be a finite probability space, and let \(\mu \) denote its associated probability mass function, assigning to each point \(\omega \in \Omega \) the real weight \(\mu (\{ \omega \} ) \ge 0\). Then there exists a point \(\omega \in \Omega \) with \(\mu (\{ \omega \} ) \gt 0\).
Let \(\Omega \) be a finite probability space, so that \(\Omega \) is finite and discrete and its canonical volume measure \(\mathbb {P}\) is a probability measure, and let \(\alpha \) be a finite type equipped with a bijection \(e : \Omega \xrightarrow {\sim } \alpha \). Then the family of singleton masses \(\bigl(\mathbb {P}(\{ e^{-1}(a)\} )\bigr)_{a \in \alpha }\) is summable with sum \(1\).
Let \(\iota \) be a finite index type and let \(\Omega \) be a finite probability space, so that the product \(\Omega ^{\iota }\) of copies of \(\Omega \) indexed by \(\iota \) carries the corresponding product probability measure. Fix an index \(i \in \iota \) and a function \(f : \Omega \to \mathbb {R}\). Then integrating the map \(\omega _{\bullet } \mapsto f(\omega _i)\) against the product measure on \(\Omega ^{\iota }\) gives the same value as integrating \(f\) against the measure on \(\Omega \):
Let \(\Omega \) be a finite probability space, and let \(p(\omega )\) denote the probability mass it assigns to each point \(\omega \in \Omega \). Then for every function \(f \colon \Omega \to \mathbb {R}\),
so that the expectation of \(f\) under the canonical probability measure equals the finite \(p\)-weighted sum of its values.
Let \(\Omega \) be a finite probability space, and let \(f : \Omega \to \mathbb {R}\) be a function satisfying \(f(\omega ) \le c\) for every \(\omega \in \Omega \), where \(c \in \mathbb {R}\) is a constant. Then the integral of \(f\) with respect to the probability measure on \(\Omega \) satisfies \(\int _\Omega f \le c\).
Let \(\Omega \) be a finite probability space and let \(f : \Omega \to \mathbb {R}\) be a function. If there is a constant \(c \in \mathbb {R}\) with \(c \lt f(\omega )\) for every \(\omega \in \Omega \), then \(c \lt \int _\Omega f\), where the integral is taken with respect to the probability measure on \(\Omega \).
Let \(\Omega \) be a finite probability space, write \(\mu \) for its canonical probability measure, and let \(p(\omega ) = \mu (\{ \omega \} )\) denote the associated probability mass function. Then for every subset \(S \subseteq \Omega \),
Let \(\Omega \) be a finite probability space and let \(S \subseteq \Omega \) be any subset. Then the probability of \(S\) equals the integral of the indicator function of \(S\) against the probability measure:
where the left-hand side is the real number obtained from the measure of \(S\).
Let \(\Omega \) be a finite probability space and let \(S \subseteq \Omega \). Then the (real-valued) probability of \(S\) equals the integral over \(\Omega \) of the indicator function of \(S\), taken with value \(1\) on \(S\) and \(0\) elsewhere:
where \(\mathbf{1}_S \colon \Omega \to \mathbb {R}\) is the function equal to \(1\) at points of \(S\) and \(0\) at points outside \(S\).
Let \(\Omega \) be a finite probability space, and let \(T\) be a finite subset of \(\Omega \). Writing \(\mathbb {P}\) for the real-valued probability measure on \(\Omega \), the measure of \(T\) is the sum of the measures of its singletons:
Let \(\Omega \) be a finite probability space and let \(f : \Omega \to \mathbb {R}\) be a nonnegative function, so that \(f(\omega ) \ge 0\) for every \(\omega \in \Omega \). Then for every real \(\varepsilon \gt 0\), the probability that \(f\) attains at least \(\varepsilon \) is bounded by the mean of \(f\) divided by \(\varepsilon \):
Let \(\Omega \) be a finite probability space and let \(\iota \) be a finite index set. If \((A_i)_{i \in \iota }\) is a pairwise disjoint family of subsets of \(\Omega \), then the probability of their union is the sum of their probabilities:
where \(\mathbb {P}(\, \cdot \, ) \in \mathbb {R}\) denotes the real-valued probability measure of \(\Omega \).
Let \(\iota \) be a finite index set, and for each \(i \in \iota \) let \(\Omega _i\) be a finite probability space carrying its canonical probability measure. For any choice of subsets \(s_i \subseteq \Omega _i\), the product measure on the finite product \(\prod _{i \in \iota } \Omega _i\) assigns to the box \(\prod _{i \in \iota } s_i\) the value
where both sides are read as real numbers.
Let \(\Xi \) be a finite probability space and let \(\Omega \) be a measurable space in which every subset is measurable. For any map \(\varphi \colon \Xi \to \Omega \) and any finite set \(T \subseteq \Omega \), the probability that \(\varphi \) lands in \(T\) decomposes as a sum over the points of \(T\):
where \(\mathrm{volume}^{\mathbb {R}}\) denotes the real-valued measure of the underlying probability measure on \(\Xi \).
Let \(\Omega _1\) and \(\Omega _2\) be finite probability spaces, with respective probability measures \(\mathbb {P}_1\) and \(\mathbb {P}_2\), and let \(A \subseteq \Omega _1\) and \(B \subseteq \Omega _2\) be any subsets. Then, under the product probability measure \(\mathbb {P}\) on \(\Omega _1 \times \Omega _2\), the measure of the rectangle \(A \times B\) factors as
where each quantity is taken as a real number.
Given a type \(\Omega \) already equipped with a MeasureSpace, Fintype, DiscreteMeasurableSpace, and IsProbabilityMeasure instance for volume, this helper packages all of those into a CommunicationComplexity.FiniteProbabilitySpace record.
Let \(\Omega _1\) and \(\Omega _2\) be finite probability spaces, and give the product \(\Omega _1 \times \Omega _2\) its induced finite probability space structure. Then for every \(x \in \Omega _1\) and \(y \in \Omega _2\), the probability mass function of the product assigns to the pair \((x, y)\) the product of the individual masses; that is, its value at \((x, y)\) equals the value of the mass function of \(\Omega _1\) at \(x\) times the value of the mass function of \(\Omega _2\) at \(y\).
Let \(\Omega \) be a finite probability space, and let \(\mu (\{ \omega \} )\) denote the probability mass that its associated probability mass function assigns to a point \(\omega \in \Omega \). Then this weight, taken as a real number, is nonnegative:
Let \(\Omega \) be a finite probability space, and for each \(\omega \in \Omega \) let \(p(\omega )\) denote the weight that the associated probability mass function assigns to \(\omega \), viewed as a real number. Then these weights sum to one:
Let \(\Omega \) be a finite probability space, and let \(f\colon \Omega \to \mathbb {R}\) be any real-valued function. Then the square of the expectation of \(f\) is at most the expectation of its square:
where both integrals are taken with respect to the underlying probability measure.
For a finite probability space \(\Omega \), toPMF converts the canonical volume measure to a probability mass function \(\mathrm{PMF}\, \Omega \), assigning to each \(\omega \) the real weight \(\mu (\{ \omega \} )\).
Let \(x \colon \mathrm{Fin}\, n \to \mathrm{Bool}\) be an \(n\)-bit Boolean input, and let \(i, j \in \mathrm{Fin}\, n\) be coordinates with \(j \ne i\). Then the input obtained from \(x\) by flipping its \(i\)-th coordinate agrees with \(x\) at coordinate \(j\); that is, its value at \(j\) equals \(x(j)\).
Let \(n \in \mathbb {N}\), let \(i \in \mathrm{Fin}\, n\) be a coordinate index, and let \(x\) be an \(n\)-bit Boolean input. The single-coordinate flip of \(x\) at \(i\) — the input obtained from \(x\) by negating its \(i\)-th coordinate and leaving all others unchanged — has value \(\neg \, x_i\) at coordinate \(i\).
Fix \(n \in \mathbb {N}\) and a coordinate \(i \in \mathrm{Fin}\, n\). On the space of \(n\)-bit Boolean inputs, functions \(\mathrm{Fin}\, n \to \mathrm{Bool}\), consider the single-coordinate flip that negates the \(i\)-th bit of its argument and leaves every other bit unchanged. This map is a bijection.
Fix \(n \in \mathbb {N}\), a coordinate \(i \in \mathrm{Fin}\, n\), and an \(n\)-bit Boolean input \(x \colon \mathrm{Fin}\, n \to \mathrm{Bool}\). Writing \(\mathrm{flip}_i\) for the operation that negates the \(i\)-th coordinate of an input and leaves all other coordinates unchanged, applying it twice recovers the original input:
Let \(n \geq 1\), and let \(\mathrm{disjointness}(n, \cdot , \cdot )\) denote the two-party Boolean function that, given subsets \(X, Y \subseteq [n]\) held by Alice and Bob respectively, returns \(\mathtt{true}\) exactly when \(X \cap Y = \emptyset \). Then the deterministic communication complexity of this function is exactly \(n + 1\).
Fix \(n \in \mathbb {N}\), and let both players’ inputs range over subsets of \([n] = \{ 0, 1, \dots , n-1\} \), so that Alice holds \(X \subseteq [n]\) and Bob holds \(Y \subseteq [n]\). The deterministic communication complexity of the set-disjointness function, which returns \(\mathtt{true}\) exactly when \(X \cap Y = \emptyset \), is at most \(n + 1\).
For every \(n \in \mathbb {N}\) and all subsets \(X, Y \subseteq [n]\), the disjointness function is symmetric in its two arguments: \(\mathrm{disjointness}(n, X, Y) = \mathrm{disjointness}(n, Y, X)\). That is, since \(X\) and \(Y\) are disjoint if and only if \(Y\) and \(X\) are disjoint, swapping the two sets leaves the Boolean value unchanged.
Fix \(n \in \mathbb {N}\), and consider the set-disjointness function on subsets of \([n]\), namely the Boolean function \(g(X, Y)\) that is \(\mathtt{true}\) exactly when \(X \cap Y = \emptyset \). Then the collection \(\{ (X, X^c) \mid X \subseteq [n]\} \) of complementary pairs is a fooling set for \(g\): for every rectangle \(R = A \times B\) (with \(A, B\) families of subsets of \([n]\)) that is monochromatic for \(g\), the intersection of \(R\) with this collection contains at most one point.
Fix an integer \(n \ge 1\), and consider the set-disjointness function on subsets of \([n]\): it takes a pair of subsets \(X, Y \subseteq [n]\) — Alice holding \(X\) and Bob holding \(Y\) — to \(\mathtt{true}\) exactly when \(X \cap Y = \emptyset \). Then the deterministic communication complexity of this function is at least \(n + 1\).
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the disjointness problem on inputs of length \(n\), drawn under the hard distribution. Write \((x_T, y_T)\) for the pair of Alice’s and Bob’s bits at the special coordinate \(T\), and for a value \(z\) of the conditioning statistic let \(\mu _z\) denote the conditional law of this bit-pair on \(\{ 0,1\} ^2\) given that the statistic equals \(z\). Suppose \(z\) is \(\gamma \)-good, in the sense that the total variation distance between \(\mu _z\) and the uniform distribution on \(\{ 0,1\} ^2\) is at most \(2\gamma \). Then for every bit-pair \(b \in \{ 0,1\} ^2\),
Fix a positive integer \(n\), and work on the sample space of the hard distribution for set disjointness, where each sample determines a special coordinate \(T \in \{ 0,\dots ,n-1\} \) together with Alice’s and Bob’s input bits. Write \(X_{\lt T}\) for Alice’s bits at the coordinates strictly below \(T\) (padded with \(\mathrm{false}\) elsewhere), \(Y_{\ge T}\) for Bob’s bits at the coordinates from \(T\) onward (padded with \(\mathrm{false}\) elsewhere), and \(Y_{\gt T}\) for Bob’s bits at the coordinates strictly above \(T\) (padded with \(\mathrm{false}\) elsewhere). Then, under the hard distribution conditioned on the event that Alice’s and Bob’s input sets are disjoint and further on Bob’s bit at the special coordinate being \(\mathrm{false}\), the triple \((T, X_{\lt T}, Y_{\ge T})\) equals the triple \((T, X_{\lt T}, Y_{\gt T})\) almost everywhere.
Let \(\omega \) be a sample of the hard distribution over \(n \ge 1\) coordinates, with special coordinate \(T\) and associated Alice/Bob bit vectors, and form its dual sample by reversing the coordinate order and interchanging the roles of Alice and Bob. Then Alice’s corrected conditioning variable of the dual sample, namely the triple \((T, X_{\lt T}, Y_{\ge T})\), equals the recoding of Bob’s corrected conditioning variable \((T, X_{\le T}, Y_{\gt T})\) of \(\omega \) that reverses the special coordinate and interchanges the two boolean vectors while reversing the coordinate order of each.
Let \(\omega \) be a sample of the hard distribution, with special coordinate \(T\), and let \(X_{\lt T}\) denote Alice’s input bits at the coordinates strictly before \(T\) (padded with \(\mathrm{false}\) elsewhere), \(Y_{\ge T}\) Bob’s input bits at the coordinates from \(T\) onward (padded with \(\mathrm{false}\) elsewhere), and \(Y_{\gt T}\) Bob’s input bits at the coordinates strictly after \(T\) (padded with \(\mathrm{false}\) elsewhere). If Bob’s bit at the special coordinate is \(\mathrm{false}\), then Alice’s corrected special-coordinate conditioning variable and the coarse conditioning variable coincide:
Fix \(n \ge 1\), and for a sample \(\omega \) from the hard distribution write \(T\) for its special coordinate, \(X_{\lt T}\) for Alice’s input bits at the coordinates before \(T\) (padded with \(\mathrm{false}\) from \(T\) onward), and \(Y_{\ge T}\) for Bob’s input bits at the coordinates from \(T\) onward (padded with \(\mathrm{false}\) before \(T\)). Then Alice’s corrected special-coordinate conditioning variable, namely the triple \((T,\, X_{\lt T},\, Y_{\ge T})\), coincides with the pair consisting of \(T\) together with Alice’s corrected dynamic conditioning data \((X_{\lt T},\, Y_{\ge T})\); that is, the two agree as functions of \(\omega \).
Let \(n\) be a positive integer and let \(p\) be a deterministic communication protocol for the disjointness problem on the \(n\) coordinates, in which each player holds a subset of \(\{ 0,\dots ,n-1\} \) and the protocol outputs a Boolean value. Running \(p\) on the input pair \((X(\omega ),Y(\omega ))\) determined by a hard sample \(\omega \) produces a transcript \(M(\omega )\). Write \(\mu \) for the hard distribution conditioned on the event that \(X(\omega )\) and \(Y(\omega )\) are disjoint, and for a sample \(\omega \) let \(T\) be its special coordinate and \(x_T\) Alice’s bit there; the dynamic conditioning of \(\omega \) is the pair \((x_{\lt T},\, y_{\ge T})\) recording Alice’s bits at coordinates below \(T\) and Bob’s bits at coordinates at or above \(T\), with every remaining coordinate set to \(\mathrm{false}\). Then for each coordinate \(i\),
where the left-hand conditional mutual information is taken under \(\mu \) further conditioned on the event \(\{ T=i\} \), and the right-hand one under \(\mu \) with the fixed-coordinate conditioning \((x_{\lt i},\, y_{\ge i})\) (Alice’s bits at coordinates below \(i\) and Bob’s bits at coordinates at or above \(i\), all other coordinates \(\mathrm{false}\)) and Alice’s bit \(x_i\) at coordinate \(i\).
Fix a length \(n \ge 1\) and a deterministic communication protocol \(p\) for the set-disjointness problem, whose inputs are subsets of \(\{ 0, \dots , n-1\} \). Write \(\mu \) for the hard-distribution measure on samples conditioned on the event that Alice’s set \(X\) and Bob’s set \(Y\) are disjoint, let \(T\) be the special coordinate, and let \(M_q\) denote the transcript produced by a protocol \(q\) on the input \((X, Y)\). Let \(p^\ast \) be the protocol dual, obtained from \(p\) by exchanging the roles of Alice and Bob and reversing the coordinate order of both input sets. Then the conditional mutual information (under \(\mu \)) between Alice’s bit \(x_T\) at the special coordinate and the transcript of \(p^\ast \), given the special coordinate together with Alice’s bits strictly before \(T\) and Bob’s bits from \(T\) onward, equals the conditional mutual information between Bob’s bit \(y_T\) at the special coordinate and the transcript of \(p\), given the special coordinate together with Alice’s bits up to and including \(T\) and Bob’s bits strictly after \(T\):
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), and let \(M\) denote the transcript that \(p\) produces on the input pair generated by a hard-distribution sample. All quantities below are computed under the hard distribution conditioned on the event that the generated pair of sets is disjoint. Writing \(X_i\) for Alice’s bit at coordinate \(i\), \(X_{\lt i}\) for Alice’s input bits at the coordinates below \(i\), \(Y_{\ge i}\) for Bob’s input bits at the coordinates \(i\) and above, and \(T\) for the special coordinate, one has
so the conditional mutual information between Alice’s bit at the special coordinate and the transcript, given the triple \((T, X_{\lt T}, Y_{\ge T})\), is the average over the \(n\) coordinates of the corresponding fixed-coordinate information terms.
Let \(n\) be a positive integer and let \(p\) be a deterministic protocol for disjointness inputs of length \(n\), and write \(M\) for the transcript produced by running \(p\) on the input pair generated by a hard-distribution sample. Work throughout under the hard distribution conditioned on the generated pair being disjoint, and write \(T\) for the special coordinate, \(X_T\) for Alice’s bit at that coordinate, and \(X_{\lt T}\) and \(Y_{\ge T}\) for Alice’s bits before the special coordinate and Bob’s bits from the special coordinate onward. Then Alice’s special-coordinate information term is the uniform average, over the \(n\) possible values of \(T\), of the corresponding conditional mutual information computed on the fiber where \(T\) is fixed:
where \(I_{\, T=i}\) denotes conditional mutual information taken under the same disjoint-conditioned distribution further conditioned on the event \(T = i\).
Let \(n\) be a positive integer and let \(p\) be a deterministic two-party communication protocol for disjointness on inputs of length \(n\), with communication complexity \(\ell \). Draw a hard sample and condition on the event \(D\) that Alice’s generated set \(X\) and Bob’s generated set \(Y\) are disjoint; let \(T\) be the special coordinate, \(X_T\) Alice’s bit at \(T\), and \(M\) the transcript obtained by running \(p\) on the generated input pair \((X,Y)\). Then, under this disjointness-conditioned hard distribution, the conditional mutual information between Alice’s special-coordinate bit and the transcript, given the special coordinate together with Alice’s bits before it and Bob’s bits from it onward, satisfies
Fix an integer \(n \ge 1\) and a deterministic two-party communication protocol \(p\) for set disjointness on inputs of length \(n\), in which Alice and Bob each hold a subset of \(\{ 1,\dots ,n\} \). Let \(\mu \) denote the hard distribution over such input pairs conditioned on the event that Alice’s and Bob’s sets are disjoint, and take all mutual information \(I(\, \cdot \, ;\, \cdot \mid \, \cdot \, )\) below with respect to \(\mu \). Under \(\mu \), let \(X, Y \colon \{ 1,\dots ,n\} \to \{ 0,1\} \) be the input vectors of Alice and Bob, let \(T\) be the special coordinate, let \(X_T\) be Alice’s bit at coordinate \(T\), and let \(M\) be the transcript produced by running \(p\) on the pair \((X,Y)\); write \(X_{\lt T}\) for the vector of Alice’s bits at coordinates strictly below \(T\), padded with \(0\) at every other coordinate, and \(Y_{\ge T}\) for the vector of Bob’s bits at coordinates at least \(T\), padded with \(0\) at every other coordinate. Then
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), and give the hard sample space the distribution conditioned on the event that the generated input pair \((X,Y)\) is disjoint. Write \(T\) for the special coordinate, \(X_T\) and \(Y_T\) for Alice’s and Bob’s bits at \(T\), and \(M\) for the transcript produced by running \(p\) on the input \((X,Y)\). Let Alice’s information term be the conditional mutual information
Bob’s information term be \(I\bigl(Y_T \, ;\, M \, \big|\, T,\, X_{\le T},\, Y_{\gt T}\bigr)\), and the total claim information be the sum of the two. Then Alice’s information term is at most the total claim information.
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), and let \(\mu \) be the hard distribution on samples, conditioned on the generated input pair \((X,Y)\) being disjoint and on Bob’s special bit \(Y_T\) being false. Write \(X_T\) for Alice’s bit at the special coordinate \(T\) and \(\Pi \) for the transcript that \(p\) produces on the generated input. Then the conditional mutual information of \(X_T\) and \(\Pi \) given the triple \((T,\, X_{\lt T},\, Y_{\ge T})\) equals the conditional mutual information of \(X_T\) and \(\Pi \) given the coarser triple \((T,\, X_{\lt T},\, Y_{\gt T})\), both computed under \(\mu \); that is,
Here \(X_{\lt T}\) denotes Alice’s bits strictly before the special coordinate (padded with false elsewhere), \(Y_{\gt T}\) Bob’s bits strictly after it, and \(Y_{\ge T}\) Bob’s bits from the special coordinate onward.
Fix a positive integer \(n\) and a hard-distribution sample \(\omega \) with special coordinate \(T \in \{ 0,\dots ,n-1\} \). Write \(\mathrm{rev}\) for the order-reversing map \(i \mapsto n-1-i\) on coordinates, and for a Boolean vector \(u\) write \(\bar u = u \circ \mathrm{rev}\) for its coordinate reversal. Recall that Alice’s conditioning variable of \(\omega \) is the triple \(\bigl(T,\ X_{\lt T},\ Y_{\ge T}\bigr)\) — the special coordinate together with Alice’s bits strictly before \(T\) (padded with \(\mathrm{false}\) elsewhere) and Bob’s bits from \(T\) onward — while Bob’s conditioning variable is \(\bigl(T,\ X_{\le T},\ Y_{\gt T}\bigr)\). Let \(\omega ^\ast \) denote the dual sample, obtained from \(\omega \) by reversing the coordinate order and interchanging Alice’s and Bob’s roles. Then Bob’s conditioning variable of \(\omega ^\ast \) equals the dual recoding of Alice’s conditioning variable of \(\omega \); that is, if \(\bigl(T,\ X_{\lt T},\ Y_{\ge T}\bigr)\) is Alice’s conditioning variable of \(\omega \), then Bob’s conditioning variable of \(\omega ^\ast \) is exactly \(\bigl(\mathrm{rev}(T),\ \overline{Y_{\ge T}},\ \overline{X_{\lt T}}\bigr)\).
Fix a positive integer \(n\) and write \(\mu \) for the hard distribution on samples \(\omega \), namely the uniform measure on the finite sample space conditioned on the disjointness event \(X(\omega ) \cap Y(\omega ) = \emptyset \); here \(X(\omega ), Y(\omega ) \subseteq \{ 0, 1, \dots , n-1\} \) are Alice’s and Bob’s input sets determined by \(\omega \), with distinguished special coordinate \(T(\omega )\) and bits \(x_T, y_T\) of Alice and Bob at that coordinate. For a deterministic two-party protocol \(p\) on these inputs, let \(M_p(\omega )\) be the transcript produced by running \(p\) on the input pair \((X(\omega ), Y(\omega ))\), and let \(p^\ast \) be its dual, obtained by exchanging the roles of Alice and Bob and reversing the coordinate order. Then, with all conditional mutual informations computed under \(\mu \),
where \(x_{\lt T}\) and \(x_{\le T}\) denote Alice’s input bits strictly before, respectively up to and including, the special coordinate (padded with \(\mathrm{false}\) at all other coordinates), and \(y_{\gt T}\) and \(y_{\ge T}\) denote Bob’s input bits strictly after, respectively from, the special coordinate onward (padded with \(\mathrm{false}\) elsewhere).
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), with communication complexity \(\ell \). Sample an input from the hard distribution conditioned on the event that Alice’s and Bob’s generated sets are disjoint; let \(M\) denote the transcript that \(p\) produces on the sampled input pair, let \(T\) denote the special coordinate, let \(Y_T\) denote Bob’s bit at the special coordinate, and let \((T, X_{\le T}, Y_{\gt T})\) denote Bob’s special-coordinate conditioning variable. Then the conditional mutual information between \(Y_T\) and \(M\) given \((T, X_{\le T}, Y_{\gt T})\) satisfies
Fix an integer \(n \ge 1\) and a deterministic communication protocol \(p\) for the disjointness problem on \(n\) coordinates, in which Alice and Bob each hold a subset of \(\{ 1,\dots ,n\} \). Work under the hard distribution over input pairs conditioned on the event that Alice’s set \(X\) and Bob’s set \(Y\) are disjoint, and let \(M\) be the transcript that \(p\) produces on a sampled input pair. Let \(T\) be the special coordinate, let \(Y_T\) be Bob’s bit at coordinate \(T\), let \(X_{\le T}\) be the vector recording Alice’s bits at all coordinates up to and including \(T\) and \(0\) beyond, let \(Y_{\gt T}\) be the vector recording Bob’s bits at all coordinates strictly greater than \(T\) and \(0\) elsewhere, and let \(X\) and \(Y\) be the full input vectors of Alice and Bob. With every mutual information taken with respect to the hard distribution conditioned on disjointness of \(X\) and \(Y\),
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the disjointness problem on inputs indexed by \(\{ 0,1,\dots ,n-1\} \). Draw a sample from the hard distribution over inputs, conditioned on the event that Alice’s set and Bob’s set are disjoint, and consider the following random variables: the special coordinate \(T\); Alice’s input bits \(X_{\le T}\) at coordinates up to and including \(T\) (padded by \(0\) elsewhere); Bob’s input bits \(Y_{\gt T}\) at coordinates strictly above \(T\) (padded by \(0\) elsewhere); Bob’s bit \(Y_T\) at the special coordinate; and the transcript \(M\) obtained by running \(p\) on the generated input pair. Then the conditional mutual information of \(Y_T\) and \(M\) given the tuple \((T, X_{\le T}, Y_{\gt T})\) is nonnegative:
Fix \(n \ge 1\), and consider the finite sample space whose points \(\omega \) each consist of a distinguished coordinate \(T \in \{ 1,\dots ,n\} \), a pair of bits \(x_T, y_T\) at that coordinate, and a labelling of every coordinate by one of the three disjoint-coordinate types (neither, left-only, right-only). From such an \(\omega \) one reads off the generated coordinate-pair vector together with the pair of input sets \(X(\omega ), Y(\omega )\), and one records the model data \(\mathrm{model}(\omega ) = (T, \text{coordinate-pair vector},\ \text{the ignored label at } T)\). Then for every triple \(z\) in the target of the model map, there is exactly one sample \(\omega \) whose generated inputs are disjoint, that is \(X(\omega )\) and \(Y(\omega )\) are disjoint, and for which \(\mathrm{model}(\omega ) = z\).
Fix a positive integer \(n\), an index \(i \in \{ 0, 1, \dots , n-1\} \), and a function \(\mathrm{coords}\) assigning to each index in \(\{ 0, 1, \dots , n-1\} \) one of the three disjointness labels *neither*, *leftOnly*, or *rightOnly*. Then there are exactly three triples \((a, f, c)\), where \(a\) is an index, \(f\) is such a labelling function, and \(c\) is a single disjointness label, subject to the constraints \(a = i\) and \(f = \mathrm{coords}\).
Fix a positive integer \(n\), a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), an achievable value \(z\) of the associated \(Z\)-variable, and a pair of bits \(b=(b_1,b_2)\). Let \(F\) be the fiber consisting of those hard-distribution samples \(\omega \) whose \(Z\)-variable equals \(z\), and for such a sample write \(x_T(\omega )\) for Alice’s bit at the special coordinate and \(y_T(\omega )\) for Bob’s bit at the special coordinate. Then
Fix a positive integer \(n\) and identify the coordinate set with \(\{ 0,1,\dots ,n-1\} \). Draw a sample \(\omega \) from the hard distribution on inputs, conditioned on the event that Alice’s set \(X(\omega )\) and Bob’s set \(Y(\omega )\) are disjoint, and write \(T\) for the special coordinate of \(\omega \) and \(X_T, Y_T\) for Alice’s and Bob’s bits at \(T\). For a deterministic protocol \(p\) whose two inputs both range over subsets of \(\{ 0,\dots ,n-1\} \), let \(\Pi _p\) be the transcript that \(p\) produces on the pair \(\bigl(X(\omega ), Y(\omega )\bigr)\). Let \(X_{\lt T}\) and \(X_{\le T}\) be the vectors of Alice’s bits at the coordinates below \(T\), respectively at the coordinates at most \(T\), each set to \(0\) outside that range, and let \(Y_{\ge T}\) and \(Y_{\gt T}\) be the vectors of Bob’s bits at the coordinates at least \(T\), respectively above \(T\), each set to \(0\) outside that range. Define the claim information of \(p\) by
both conditional mutual informations being taken under the conditioned distribution. Let \(p^{\dagger }\) be the protocol obtained from \(p\) by exchanging the roles of Alice and Bob throughout and reversing the coordinate order \(i \mapsto n-1-i\) on both parties’ input sets. Then \(\mathcal{I}(p^{\dagger }) = \mathcal{I}(p)\).
Fix a positive integer \(n\) and let \(p\) be a deterministic two-party communication protocol whose inputs are the pairs of subsets of \(\{ 1,\dots ,n\} \) held by Alice and Bob in the set-disjointness problem. Draw a sample from the hard distribution on such inputs conditioned on the event that Alice’s set \(X\) and Bob’s set \(Y\) are disjoint, let \(M\) be the transcript that \(p\) produces on the resulting input pair, let \(T\) be the special coordinate carried by the sample, and write \(X_T\) and \(Y_T\) for Alice’s and Bob’s bits at \(T\); for a threshold coordinate, write \(X_{\lt T}\), \(X_{\le T}\), \(Y_{\gt T}\), \(Y_{\ge T}\) for the tuples of input bits at the coordinates strictly before, up to and including, strictly after, and from \(T\) onward. Then, writing \(C(p)\) for the communication complexity of \(p\) (the worst-case number of bits exchanged) and \(\ln \) for the natural logarithm, the sum of the conditional mutual information of \(X_T\) with \(M\) given \((T, X_{\lt T}, Y_{\ge T})\) and the conditional mutual information of \(Y_T\) with \(M\) given \((T, X_{\le T}, Y_{\gt T})\) satisfies
Fix \(n \ge 1\), and let \(\omega \) be a sample of the hard distribution: a special coordinate \(T \in \{ 0,1,\dots ,n-1\} \), bits \(x_T\) and \(y_T\) for Alice and Bob at that coordinate, and a disjointness type assigned to every coordinate. From \(\omega \) form the coarse conditioning \((T,\, X_{\lt T},\, Y_{\gt T})\), where \(X_{\lt T}\) is Alice’s input bits at the coordinates strictly below \(T\) (padded with \(\mathrm{false}\) elsewhere) and \(Y_{\gt T}\) is Bob’s input bits at the coordinates strictly above \(T\) (padded with \(\mathrm{false}\) elsewhere). Let \(\omega '\) be obtained from \(\omega \) by toggling only Alice’s special bit \(x_T\), leaving \(T\), \(y_T\), and all coordinate types unchanged. Then \(\omega \) and \(\omega '\) produce the same coarse conditioning: \((T,\, X_{\lt T},\, Y_{\gt T})\) is identical for \(\omega '\) and \(\omega \).
Fix a deterministic communication protocol \(p\) for the disjointness problem on \(n\) coordinates and an achievable value \(z\) of the summary variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), whose components are the transcript \(M\), the special coordinate \(T\), the vector \(X_{\lt T}\) of Alice’s bits at coordinates below \(T\), and the vector \(Y_{\gt T}\) of Bob’s bits at coordinates above \(T\). Let \(\nu \) be the uniform distribution on hard samples conditioned on the fiber \(\{ Z = z\} \), let \(\mu \) be the law of the special bit-pair \((X_T, Y_T)\) under \(\nu \), and let \(\mu _X\) and \(\mu _Y\) be the laws of the single bits \(X_T\) and \(Y_T\) under \(\nu \). Then
Fix a positive integer \(n\), a deterministic disjointness protocol \(p\) on inputs of length \(n\), and an achievable value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\). Write \(F\) for the fiber of \(Z\) above \(z\), and for a subset \(S\) of the sample space let \(\mathrm{vol}(S)\) denote its real-valued volume. Suppose that for every pair of bits \((b_1, b_2) \in \mathrm{Bool} \times \mathrm{Bool}\),
Then the conditional law of the special bit-pair on the fiber \(F\) equals the product of the conditional law of Alice’s special bit and the conditional law of Bob’s special bit.
Fix a positive integer \(n\) and a deterministic disjointness protocol \(p\) on inputs of length \(n\), and let \(z\) be an achievable value of the variable \(Z=(M,T,X_{\lt T},Y_{\gt T})\). Write \(\mu \) for the conditional law, on the fiber \(\{ Z=z\} \), of the special bit-pair \((X_T,Y_T)\in \{ 0,1\} ^2\), and write \(\mu _X\) and \(\mu _Y\) for the conditional laws on that same fiber of the individual special bits \(X_T\) and \(Y_T\). Suppose that for every pair \(b=(b_1,b_2)\in \{ 0,1\} ^2\) the mass assigned by \(\mu \) to the singleton \(\{ b\} \) equals the product \(\mu _X(\{ b_1\} )\, \mu _Y(\{ b_2\} )\). Then \(\mu \) is the product measure \(\mu _X\otimes \mu _Y\).
Fix a deterministic disjointness protocol \(p\) on inputs of length \(n\) and an achievable value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), and let \(\mu _z\) denote the law of the hard-distribution sample conditioned on the fiber \(\{ Z = z\} \). Suppose that for every pair of bits \((b_1, b_2) \in \mathrm{Bool} \times \mathrm{Bool}\) the \(\mu _z\)-probability that the special-coordinate bits of Alice and Bob equal \((b_1, b_2)\) is the product of the \(\mu _z\)-probability that Alice’s special bit is \(b_1\) and the \(\mu _z\)-probability that Bob’s special bit is \(b_2\). Then the conditional law of the special bit-pair under \(\mu _z\) equals the product of the conditional law of Alice’s special bit and the conditional law of Bob’s special bit.
Fix a positive length \(n\), a deterministic disjointness protocol \(p\) on inputs of length \(n\), and an achievable value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\). Let \(\mu _z\) be the conditional distribution of a hard-distribution sample given \(Z = z\), and for a sample \(\omega \) write its special pair as \((x_T(\omega ), y_T(\omega ))\), the two players’ bits at the special coordinate \(T = T(\omega )\). Then for every bit-pair \(b \in \{ 0,1\} \times \{ 0,1\} \), the conditional law of the special pair — the pushforward of \(\mu _z\) along the special-pair map — assigns to the singleton \(\{ b\} \) exactly the \(\mu _z\)-probability of the event that the special pair equals \(b\); that is, its real-valued mass at \(\{ b\} \) equals the \(\mu _z\)-measure of the set of samples \(\omega \) with \((x_T(\omega ), y_T(\omega )) = b\).
Fix a deterministic disjointness protocol \(p\) on inputs of length \(n\) and an achievable value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\). Write \(\mu \) for the conditional law of the special-coordinate bit-pair on the fiber \(\{ Z = z\} \), and \(\mu _X\), \(\mu _Y\) for the conditional laws on that same fiber of Alice’s bit and of Bob’s bit at the special coordinate, all measures on Boolean values. Suppose the joint law factors as the product of its marginals, \(\mu = \mu _X \otimes \mu _Y\). Then for every pair of bits \(b = (b_1, b_2) \in \mathrm{Bool} \times \mathrm{Bool}\), the real-valued mass assigned to the singleton is the product of the marginal masses,
Fix a positive integer \(n\) and consider deterministic communication protocols whose two players, Alice and Bob, each receive a subset of an \(n\)-element coordinate set. The hard distribution is the uniform law on a finite sample space in which a sample fixes a distinguished coordinate \(T\) together with both players’ bits at every coordinate, and hence an input pair. Given a protocol \(q\), define the summary of a sample to be the four-tuple \((\tau , t, a, b)\), where \(\tau \) is the transcript \(q\) produces on the induced input pair, \(t = T\), the vector \(a\) gives Alice’s bits at the coordinates strictly below \(T\) and \(\mathrm{false}\) at every other coordinate, and the vector \(b\) gives Bob’s bits at the coordinates strictly above \(T\) and \(\mathrm{false}\) at every other coordinate.
Let \(p\) be such a protocol and let \(z = (\tau , t, a, b)\) be a summary realised by at least one sample of \(p\). Let \(p^{\ast }\) be the dual protocol, formed by exchanging the roles of Alice and Bob and reversing the coordinate order of both input sets via \(i \mapsto n+1-i\). Let \(z^{\ast }\) be the summary of \(p^{\ast }\) whose special coordinate is the reversal \(n+1-t\), whose Alice-vector is the coordinate reversal of \(b\), whose Bob-vector is the coordinate reversal of \(a\), and whose transcript is the image of \(\tau \) under the map on transcripts induced by passing to \(p^{\ast }\).
Then the law of Alice’s bit at the special coordinate, for the hard distribution conditioned on the event that the summary under \(p^{\ast }\) equals \(z^{\ast }\), coincides with the law of Bob’s bit at the special coordinate, for the hard distribution conditioned on the event that the summary under \(p\) equals \(z\).
Fix a deterministic disjointness protocol \(p\) on inputs of length \(n\) and an achievable value \(z\) of the record \(Z=(M,T,X_{\lt T},Y_{\gt T})\), and write \(\mu _z\) for the hard distribution conditioned on the fiber \(Z=z\). Let \(X_T\) and \(Y_T\) denote Alice’s and Bob’s bits at the special coordinate \(T\), and suppose that under \(\mu _z\) the pair \((X_T,Y_T)\) has product law, i.e. the law of \((X_T,Y_T)\) equals the product of the law of \(X_T\) and the law of \(Y_T\). Then for every \(b_Y\in \{ \text{false},\text{true}\} \) such that the event \(\{ Y_T=b_Y\} \) has positive \(\mu _z\)-mass, the law of \(X_T\) under \(\mu _z\) coincides with the law of \(X_T\) under \(\mu _z\) conditioned further on \(\{ Y_T=b_Y\} \).
Fix a positive integer \(n\), a deterministic disjointness protocol \(p\) on inputs of length \(n\), and an achievable value \(z\) of the coarse variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\). Let \(\mu _z\) be the hard distribution conditioned on the fiber \(\{ Z = z\} \), and suppose that under \(\mu _z\) the event that Bob’s bit at the special coordinate is \(\mathrm{false}\) (that is, \(Y_T = \mathrm{false}\)) has positive probability. Then the law of Alice’s special bit \(X_T\) under \(\mu _z\) coincides with the law of \(X_T\) under \(\mu _z\) conditioned further on \(\{ Y_T = \mathrm{false}\} \).
Fix a positive integer \(n\), a deterministic protocol \(p\) for length-\(n\) set disjointness, a value \(z\) in the range of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), and a bit \(b\). Let \(\mu _z\) denote the hard input distribution conditioned on the fiber \(\{ Z = z\} \), and let \(X_T\) denote Alice’s bit at the special coordinate \(T\). Then the conditional law of \(X_T\) on this fiber — the pushforward of \(\mu _z\) along the map \(X_T\) — assigns to the singleton \(\{ b\} \) exactly the \(\mu _z\)-probability of the event that Alice’s special bit equals \(b\); that is, its real-valued mass at \(b\) equals \(\mu _z\big(\{ \, X_T = b\, \} \big)\).
Let \(n\) be a positive integer, let \(p\) be a deterministic protocol for disjointness inputs of length \(n\), let \(z\) be an achievable value of the variable \(Z=(M,T,X_{\lt T},Y_{\gt T})\) for \(p\), and let \(b\) be a Boolean value. Write \(\mu _z\) for the hard-distribution measure conditioned on the fiber \(\{ Z=z\} \), and let \(S\) be the map sending a sample to Bob’s bit at its special coordinate. Taking the conditional law of Bob’s special bit to be the pushforward \(S_*\mu _z\), its real-valued mass on the singleton \(\{ b\} \) equals the real-valued mass that \(\mu _z\) assigns to the preimage \(S^{-1}(\{ b\} )\).
Fix a length \(n \ge 1\) and a deterministic two-party protocol \(p\) for the disjointness problem on inputs of that length, and equip the hard-distribution sample space with the law obtained by conditioning on the generated input pair being disjoint and, in addition, on Bob’s bit \(Y_T\) at the special coordinate \(T\) being \(\mathrm{false}\). Write \(X_T\) for Alice’s bit at the special coordinate, and let \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) be the random variable that records the protocol transcript \(M\) produced by running \(p\), the special coordinate \(T\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\). Then the average over \(Z\) of the Kullback–Leibler divergence between the conditional law of \(X_T\) given \(Z\) and the uniform distribution on a single bit equals the mutual information \(I(X_T \, ; Z)\).
Let \(p\) be a protocol on the hard sample space, let \(X, X' : \Omega \to \{ 0,1\} \) be two Boolean variables and let \(Z, Z' : \Omega \to [n] \times \{ 0,1\} ^n \times \{ 0,1\} ^n\) be two conditioning variables on the hard sample \(\Omega = \texttt{HardSample}\, n\). Suppose that \(X'\) and \(Z'\) evaluated on the dual sample are \(X\) and the recoded \(Z\) on the original sample, i.e. \(X'(\sigma (\omega )) = X(\omega )\) and \(Z'(\sigma (\omega )) = \delta (Z(\omega ))\) for every \(\omega \), where \(\sigma \) is the hard-sample duality dualHardSample and \(\delta \) the conditioning recoding dualConditioningValue. Then, under the disjoint-conditioned measure \(\mu _{\mathcal D}\), the conditional mutual information between \(X'\) and the transcript of the dual protocol \(\bar p\) given \(Z'\) equals that between \(X\) and the transcript of \(p\) given \(Z\):
where \(M_p = \texttt{message}\, n\, p\) denotes the transcript of \(p\).
Fix a positive integer \(n\), and let \(\mu _n\) be the hard input distribution on pairs of subsets \((X, Y)\) of \([n] = \{ 0, 1, \dots , n-1\} \). Let \(p\) be a deterministic two-party communication protocol for set disjointness on inputs from \([n]\), and suppose its distributional error against the disjointness function under \(\mu _n\) is at most \(\tfrac {1}{32}\), i.e. the \(\mu _n\)-probability that the output of \(p\) disagrees with \(\mathrm{disjointness}(n, X, Y)\) is at most \(\tfrac {1}{32}\). Then the communication complexity of \(p\) satisfies
Fix \(n \ge 1\), an index \(i \in \{ 0, 1, \dots , n-1\} \), and a target bit-string \(b \colon \{ 0,1,\dots ,n-1\} \to \{ \mathrm{false},\mathrm{true}\} \). For a configuration \(c\) that assigns to each coordinate one of the three disjointness states, write \(A_i(c) \in \{ \mathrm{false},\mathrm{true}\} ^n\) for Alice’s fixed conditioning string and \(A_{i,k}(c)\) for its value at coordinate \(k\). Then the set of configurations whose conditioning string equals \(b\) exactly is the intersection over all coordinates of the sets on which each individual conditioning bit is correct:
Fix a positive integer \(n\) and a coordinate index \(i \in \{ 0, 1, \dots , n-1\} \), and consider any assignment \(c = (c_0, \dots , c_{n-1})\) of a disjointness value to each coordinate, where each \(c_k\) carries Alice’s bit \(x_k\) and Bob’s bit \(y_k\). Then Alice’s fixed-coordinate conditioning at \(i\) — the pair whose first component holds \(x_j\) for \(j \lt i\) and \(\mathrm{false}\) elsewhere, and whose second component holds \(y_j\) for \(j \ge i\) and \(\mathrm{false}\) elsewhere — coincides with the result of first forming the single bit vector \(w\) with \(w_k = x_k\) for \(k \lt i\) and \(w_k = y_k\) for \(k \ge i\), and then recoding \(w\) into the padded pair \(\bigl(\, j \mapsto w_j \text{ if } j \lt i \text{ else } \mathrm{false},\ \ j \mapsto w_j \text{ if } i \le j \text{ else } \mathrm{false}\, \bigr)\). The two constructions agree as functions of the coordinate assignment.
Fix \(n \ge 1\) and an index \(i \in \{ 0,\dots ,n-1\} \), and consider the map that sends a bit-vector \(b \colon \{ 0,\dots ,n-1\} \to \{ 0,1\} \) to the pair of bit-vectors \((u, v)\) obtained by splitting \(b\) at coordinate \(i\): namely \(u(j) = b(j)\) for \(j \lt i\) and \(u(j) = 0\) otherwise, while \(v(j) = b(j)\) for \(j \ge i\) and \(v(j) = 0\) otherwise. This map is injective.
Let \(x, y \in \{ \mathrm{true}, \mathrm{false}\} \) be a pair of bits that are not both \(\mathrm{true}\). Then the disjoint coordinate built from \((x, y)\) has Alice’s bit equal to \(x\); that is, extracting Alice’s bit from this coordinate returns \(x\) exactly.
Let a *disjointness coordinate* be one of three values, recording, for a single index, whether that index belongs to neither of two subsets, to the first subset only (Alice’s), or to the second subset only (Bob’s). To each coordinate \(c\) attach two Boolean bits: Alice’s bit \(x(c)\), which is set precisely when \(c\) marks the index as belonging to Alice’s subset only, and Bob’s bit \(y(c)\), which is set precisely when \(c\) marks it as belonging to Bob’s subset only. Conversely, from a pair of bits \((x, y)\) form a coordinate \(\kappa (x, y)\) by taking "Alice only" if \(x\) is set, otherwise "Bob only" if \(y\) is set, and otherwise "neither." Then \(\kappa \bigl(x(c), y(c)\bigr) = c\) for every coordinate \(c\); that is, splitting a coordinate into its two bits and reassembling recovers the original coordinate.
Let \(x, y \in \{ \mathrm{true}, \mathrm{false}\} \) be Alice’s and Bob’s bits in a single coordinate, and suppose they are not both \(\mathrm{true}\). Then Bob’s bit extracted from the disjoint coordinate assembled from \(x\) and \(y\) equals \(y\).
Each coordinate of a disjointness instance takes one of three values — belonging to neither player, to the left player (Alice) only, or to the right player (Bob) only — and one records Alice’s bit, equal to \(1\) exactly on the left-only value, and Bob’s bit, equal to \(1\) exactly on the right-only value. Fix a positive integer \(n\), indices \(i, k \in \{ 0, 1, \dots , n-1\} \), and a bit \(b \in \{ 0,1\} \), and form the coordinate value that is left-only if \(k \lt i\) and \(b = 1\), right-only if \(k \ge i\) and \(b = 1\), and neither in every case where \(b = 0\). Then reading Alice’s bit of this value when \(k \lt i\), and Bob’s bit of it when \(k \ge i\), returns \(b\).
Let \(\omega \) be a hard-distribution sample, with distinguished coordinate \(T\) carrying Alice’s bit \(x_T\) and Bob’s bit \(y_T\), and let \(X(\omega )\) and \(Y(\omega )\) be Alice’s and Bob’s input sets under \(\omega \). Then \(X(\omega )\) and \(Y(\omega )\) are disjoint if and only if it is not the case that \(x_T\) and \(y_T\) are both \(\mathrm{true}\).
Consider the hard distribution on samples \(\omega \), each of which determines Alice’s input set \(X(\omega )\) and Bob’s input set \(Y(\omega )\), and let \(D\) be the disjointness event \(\{ \omega : X(\omega ) \cap Y(\omega ) = \varnothing \} \). Under the distribution obtained by conditioning the hard distribution on \(D\), almost every sample belongs to \(D\); that is, with respect to the conditioned measure, the sets \(X(\omega )\) and \(Y(\omega )\) are disjoint almost surely.
Fix a positive integer \(n\), and consider the finite sample space of hard samples \(\omega = (T, x_T, y_T, \text{other})\), where \(T \in \{ 0, \dots , n-1\} \), the bits \(x_T, y_T\) lie in \(\{ \text{false}, \text{true}\} \), and \(\text{other}\) assigns to each coordinate one of \(\{ \text{neither}, \text{leftOnly}, \text{rightOnly}\} \). Let the dual map send \(\omega \) to the sample whose distinguished coordinate is the reversal \(\mathrm{rev}(T) = n-1-T\), whose distinguished bits are \(x_T\) and \(y_T\) interchanged, and whose remaining coordinate is \(i \mapsto \mathrm{swap}\big(\text{other}(\mathrm{rev}(i))\big)\); and let the disjoint-conditioned measure be the uniform distribution on this sample space conditioned on the event that Alice’s and Bob’s induced input sets are disjoint. Then the dual map is measure-preserving for the disjoint-conditioned measure: it is measurable, and the pushforward of the disjoint-conditioned measure along it is again the disjoint-conditioned measure.
Let \(n \ge 1\), and write \(\mathcal{C} = \{ \text{neither}, \text{left-only}, \text{right-only}\} \) for the three-element type of disjoint coordinates. Consider the hard distribution on samples conditioned on the event that the generated input pair of sets is disjoint. For every point \(z\) of the product space \(\{ 1,\dots ,n\} \times \bigl(\{ 1,\dots ,n\} \to \mathcal{C}\bigr) \times \mathcal{C}\), the conditional probability that a sample’s disjoint model — namely the triple consisting of its special coordinate, its generated disjoint coordinate-pair vector, and the ignored value at the special coordinate — equals \(z\) is exactly
Since this value is independent of \(z\) and equals the reciprocal of the cardinality of the product space, the disjoint model pushes the conditioned distribution forward to the uniform distribution on that space.
Fix a positive integer \(n\), and consider the hard distribution on hard samples conditioned on the event that the generated pair \((X,Y)\) is disjoint. For every hard sample \(\omega \), this conditioned distribution assigns the same mass to the singleton \(\{ \omega ^{*}\} \) as to the singleton \(\{ \omega \} \), where \(\omega ^{*}\) is the dual sample obtained by swapping Alice’s and Bob’s roles and reversing the order of the coordinates. That is, the conditioned probability of \(\omega ^{*}\) equals that of \(\omega \).
Consider the hard distribution on samples \(\omega \), conditioned on the event that Alice’s set and Bob’s set are disjoint. Let \(b_x, b_y \in \{ 0,1\} \) be a pair of prescribed values that are not both \(1\). Then the probability, under this conditioned distribution, that the special coordinate \(T\) carries Alice’s bit equal to \(b_x\) and Bob’s bit equal to \(b_y\) is exactly \(\tfrac {1}{3}\).
Fix a positive integer \(n\), and let \(\mathcal{D}\) be the three-element set of disjoint coordinate types, each coordinate belonging to neither player’s set, to Alice’s set only, or to Bob’s set only. Under the hard distribution conditioned on the generated input pair being disjoint, the joint law of the special coordinate together with the generated disjoint coordinate-pair vector is uniform: for every index \(i \in \{ 1, \dots , n\} \) and every assignment \(c \colon \{ 1, \dots , n\} \to \mathcal{D}\), the probability that the special coordinate equals \(i\) and the generated coordinate-pair vector equals \(c\) is exactly
Fix a positive integer \(n\), and consider the hard distribution on the sample space conditioned on the event that the generated input pair \((X,Y)\) is disjoint. Writing \(T \in \{ 0,1,\dots ,n-1\} \) for the special coordinate of a sample, then for every index \(i \in \{ 0,1,\dots ,n-1\} \) the conditional probability that \(T = i\) equals \(\tfrac {1}{n}\); that is, the special coordinate remains uniformly distributed after conditioning on disjointness.
Fix \(n \ge 1\), and consider the hard distribution on \(\mathrm{HardSample}(n)\) conditioned on the disjointness event (that Alice’s set \(X\) and Bob’s set \(Y\) are disjoint). Then for every coordinate \(i \in \{ 0,1,\dots ,n-1\} \), the conditional probability that the special coordinate \(T\) equals \(i\) is exactly
In other words, conditioning on disjointness leaves the special coordinate \(T\) uniformly distributed over the \(n\) coordinates.
Fix a positive integer \(n\), and draw a point of the hard-distribution sample space uniformly at random, then condition on the event that Alice’s and Bob’s generated input sets are disjoint. Under this conditional distribution, the probability that Bob’s bit at the special coordinate \(T\) equals \(\mathrm{false}\) is exactly \(2/3\).
Fix \(n \ge 1\), and consider a sample drawn uniformly from the finite space consisting of a distinguished coordinate \(T \in \{ 0,\dots ,n-1\} \), two special bits \(x_T, y_T \in \{ 0,1\} \), and, for every coordinate, an independent label from \(\{ \text{neither},\text{leftOnly},\text{rightOnly}\} \); these data determine Alice’s set \(X = \{ i : x_i = 1\} \) and Bob’s set \(Y = \{ i : y_i = 1\} \), where at the coordinate \(T\) the two bits are \(x_T\) and \(y_T\) and at every other coordinate they are read off the label (with \(\text{neither}\mapsto (0,0)\), \(\text{leftOnly}\mapsto (1,0)\), \(\text{rightOnly}\mapsto (0,1)\)). Conditioning this uniform measure on the event that \(X\) and \(Y\) are disjoint, the conditional probability of the event \(x_T = 0\) and \(y_T = 0\) is exactly \(\tfrac {1}{3}\).
The inductive type with three constructors neither, leftOnly, rightOnly, encoding the three disjoint coordinate pairs \((A_i, B_i) \in \{ (0,0), (1,0), (0,1)\} \) placed at every coordinate \(i\) other than the special coordinate \(T\) in the hard distribution for disjointness.
Consider the type of possible roles a single coordinate can play in the disjointness communication problem: it belongs to neither of the two sets, to the left set only, or to the right set only. This is a finite type, and it has exactly three elements.
Consider a coordinate in the disjointness distribution, whose state is one of three cases — belonging to neither player, to Alice only, or to Bob only — and let \(\mathrm{xBit}(c)\) and \(\mathrm{yBit}(c)\) denote Alice’s and Bob’s bits in that coordinate. Then for every such coordinate \(c\) it is never the case that both \(\mathrm{xBit}(c)\) and \(\mathrm{yBit}(c)\) equal true.
A disjoint coordinate records, for a single index, which of the two players’ inputs contains that index: neither, Alice only, or Bob only. Writing \(\mathrm{swap}\) for the operation that exchanges the two players’ sides — fixing "neither" and interchanging "Alice only" and "Bob only" — one has \(\mathrm{swap}(\mathrm{swap}(c)) = c\) for every disjoint coordinate \(c\); that is, \(\mathrm{swap}\) is an involution.
Let \(\omega \) be a hard-distribution sample whose induced input pair, Alice’s set \(X(\omega )\) and Bob’s set \(Y(\omega )\), are disjoint. Then for every coordinate \(i\), Alice’s bit read off from the disjoint coordinate-pair generated at \(i\) agrees with Alice’s bit at \(i\) under \(\omega \).
Let \(\omega \) be a sample of the hard distribution lying in the disjoint event — that is, one whose induced input sets \(X(\omega )\) and \(Y(\omega )\) are disjoint. Then for every coordinate \(i\), the Bob-bit of the disjoint coordinate assigned to \(i\) by the generated coordinate-pair vector agrees with Bob’s bit \(\mathrm{yBit}(\omega , i)\) at that coordinate.
Let \(\omega \) be a sample of the hard distribution, and let \(\omega ^\ast \) be its dual, obtained by exchanging Alice’s and Bob’s roles while reversing the order of the coordinates. Then Alice’s and Bob’s input sets \(X(\omega ^\ast )\) and \(Y(\omega ^\ast )\) are disjoint if and only if \(X(\omega )\) and \(Y(\omega )\) are disjoint; that is, \(\omega ^\ast \) lies in the disjointness event exactly when \(\omega \) does.
Fix an integer \(n \ge 1\), and consider the finite space of hard-distribution samples on \(n\) coordinates, each sample consisting of a special coordinate \(T \in \{ 1,\dots ,n\} \), two bits \(x_T\) and \(y_T\), and a disjoint-coordinate label (one of *neither*, *left only*, *right only*) attached to every coordinate. From such a sample one forms Alice’s set \(X\) and Bob’s set \(Y\). Then, as subsets of this sample space, the event that \(X\) and \(Y\) are disjoint equals the complement of the special-intersection event; equivalently, \(X\) and \(Y\) fail to be disjoint precisely on those samples for which \(x_T\) and \(y_T\) are both true.
Fix a positive integer \(n\), and take the hard distribution on samples, conditioned first on the event that Alice’s and Bob’s input sets are disjoint and then further conditioned on the event that Bob’s special bit \(y_T\) is false. Under this doubly-conditioned distribution, Bob’s special bit \(y_T\) equals \(\mathtt{false}\) almost surely.
Fix an input length \(n \ge 1\), a deterministic disjointness protocol \(p\) on inputs of length \(n\), and an achievable value \(z\) of the summary variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\). Let \(D\) denote the hard distribution conditioned on the event that the two generated input sets are disjoint, and write \(Y_T\) for Bob’s bit at the special coordinate. Then conditioning \(D\) first on the event \(\{ Y_T = \mathrm{false}\} \) and then on the fiber \(\{ Z = z\} \) produces exactly the same measure as conditioning \(D\) in a single step on the intersection \(\{ Y_T = \mathrm{false}\} \cap \{ Z = z\} \).
Fix \(n \ge 1\), a deterministic communication protocol \(p\) for disjointness inputs on \(n\) coordinates, and a real number \(\gamma \). Let \(\mu \) be the uniform distribution on hard samples, conditioned on the event that the generated sets \(X(\omega )\) and \(Y(\omega )\) are disjoint and that Bob’s bit at the special coordinate is \(\mathrm{false}\). For a sample \(\omega \), let \(Z(\omega )\) record the transcript of \(p\) on the input pair together with the special coordinate, Alice’s input bits before it, and Bob’s input bits after it, and write
the total variation distance between the conditional law of Alice’s bit at the special coordinate given \(Z(\omega )\) and the uniform law \(U\) on one bit. If \(\int \Delta (\omega )^2 \, d\mu (\omega ) \le \gamma ^4\), then \(\int \Delta (\omega ) \, d\mu (\omega ) \le \gamma ^2\).
Fix a positive integer \(n\), and work on the finite sample space of hard samples on \(n\) coordinates, equipped with its ambient counting measure. Let \(\mu \) be the measure obtained by conditioning this counting measure first on the disjointness event—that Alice’s input set and Bob’s input set are disjoint—and then on the event that Bob’s bit at the special coordinate equals \(\mathrm{false}\). Then \(\mu \) is a probability measure; equivalently, \(\mu \) assigns total mass \(1\) to the whole sample space.
Fix a positive integer \(n\), and let \(\mu \) denote the hard distribution on samples conditioned on the event that Alice’s and Bob’s generated input sets are disjoint and further conditioned on Bob’s special bit being \(\mathsf{false}\). Write \(X_T\) for Alice’s bit at the special coordinate and write \((T, X_{\lt T}, Y_{\gt T})\) for the coarse variable recording the special coordinate together with Alice’s bits strictly before it and Bob’s bits strictly after it. Then for every bit \(b \in \{ \mathsf{false}, \mathsf{true}\} \) and every value \(c\) of the coarse variable,
where the measures are taken as real numbers; that is, under \(\mu \), Alice’s special bit is uniformly distributed over each fiber of the coarse variable.
Fix a positive integer \(n\) and a Boolean value \(b\), and work on the finite sample space of hard-distribution samples over \(n\) coordinates. Form the probability distribution obtained by taking the uniform hard distribution, conditioning on the disjointness event that Alice’s set \(X\) and Bob’s set \(Y\) are disjoint, and conditioning again on Bob’s special bit \(y_T\) taking the value \(\mathrm{false}\). Under this distribution, the probability that Alice’s special bit \(x_T\) equals \(b\) is exactly \(\tfrac {1}{2}\).
Fix a positive length \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), and let \(z\) be an achievable value of the variable \(Z=(M,T,X_{\lt T},Y_{\gt T})\) for \(p\), with fiber \(Z^{-1}(z)\). Write \(D\) for the hard distribution conditioned on the generated input pair being disjoint, and let \(D'\) be \(D\) conditioned further on the event \(\{ Y_T=\mathrm{false}\} \) that Bob’s bit at the special coordinate \(T\) is \(\mathrm{false}\). Then
Fix an integer \(n \ge 1\), and let the hard distribution be the uniform law on samples that each specify a special coordinate \(T \in \{ 1,\dots ,n\} \), Alice’s and Bob’s bits \(x_T, y_T\) at that coordinate, and, at every coordinate, a disjoint coordinate-type recording which of the two players (if either) occupies it. Conditioning the hard distribution on the event that the induced input sets of Alice and Bob are disjoint and on Bob’s special bit satisfying \(y_T = \mathrm{false}\), Alice’s special bit \(x_T\) becomes uniformly distributed on \(\{ \mathrm{true}, \mathrm{false}\} \).
Fix a positive integer \(n\), a deterministic protocol \(p\) for the disjointness function on inputs of length \(n\), and a real number \(\gamma \gt 0\). Let \(\mu \) denote the hard distribution on samples, conditioned on the generated input pair being disjoint and on Bob’s bit at the special coordinate being \(\mathrm{false}\). For a sample \(\omega \), let \(Z(\omega )\) record the transcript produced by running \(p\) on the input pair generated by \(\omega \), together with the special coordinate \(T\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\); and let
be the total variation distance between the conditional law of Alice’s bit \(x_T\) at the special coordinate, given the value \(Z(\omega )\), and the uniform distribution on a single bit. If the expectation of \(D\) under \(\mu \) satisfies \(\int D \, d\mu \le \gamma ^2\), then the \(\mu \)-probability that \(D(\omega ) \gt \gamma \) is at most \(\gamma \).
Let \(n\) be a positive integer and let \(p\) be a deterministic two-party communication protocol taking a pair of subsets of \([n]\) to a Boolean value. Then the distributional error of \(p\) against the disjointness function, measured under the hard input distribution on pairs \((X,Y)\) of subsets of \([n]\), equals the probability, taken over the hard sample space, of the event that \(p\) errs — that is, the total mass of those samples \(\omega \) for which \(p\), run on the input pair \((X(\omega ), Y(\omega ))\) generated by \(\omega \), disagrees with \(\mathrm{disjointness}(n, X(\omega ), Y(\omega ))\).
Fix a positive integer \(n\), and let a triple \(c = (i, x, y)\) consist of a coordinate \(i \in \{ 0, 1, \dots , n-1\} \) together with two Boolean vectors \(x, y \colon \{ 0, 1, \dots , n-1\} \to \{ \text{true}, \text{false}\} \). Applying the dual conditioning-value recoding — which sends \(c\) to \((n-1-i,\, y \circ \mathrm{rev},\, x \circ \mathrm{rev})\), where \(\mathrm{rev}\) reverses the coordinate order — twice in succession returns the original triple \(c\); that is, the recoding is an involution.
Fix an integer \(n \ge 1\), and for a boolean vector \(x \colon \{ 0,1,\dots ,n-1\} \to \{ \text{true},\text{false}\} \) write \(\bar x\) for its coordinate reversal, \(\bar x(i) = x(n-1-i)\). Consider the map on triples \((i, x, y)\), where \(i \in \{ 0,1,\dots ,n-1\} \) and \(x, y\) are boolean vectors of length \(n\), defined by
which reverses the index, swaps the two vectors, and reverses the coordinate order of each. This map is injective.
Fix \(n \ge 1\), and let \(\omega \) be any hard sample: a distinguished coordinate \(T \in \{ 0,\dots ,n-1\} \), two bits recording Alice’s and Bob’s values there, and an assignment of a disjoint coordinate (neither, left-only, or right-only) to every coordinate. Applying the dualization map to \(\omega \) twice returns \(\omega \) itself; that is, the dualization — which reverses the coordinate order, interchanges Alice’s and Bob’s bits at the distinguished coordinate, and swaps the two sides of every coordinate assignment — is an involution.
Fix a positive integer \(n\), and let \(D\) denote the dualization map on the finite space of hard samples, which sends a hard sample to the one obtained by exchanging Alice’s and Bob’s bits and reversing the order of the coordinates. Then for every hard sample \(\omega \), the preimage of the singleton \(\{ \omega \} \) under \(D\) is exactly the singleton \(\{ D(\omega )\} \); that is, \(D^{-1}(\{ \omega \} ) = \{ D(\omega )\} \).
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the disjointness problem on inputs of length \(n\). For a hard-distribution sample \(\omega \), let \(Z_p(\omega )\) be its \(Z\)-value — the transcript of \(p\) on the input pair induced by \(\omega \), together with the special coordinate and the input bits lying strictly before and strictly after it — and let \(\alpha (\omega )\) and \(\beta (\omega )\) denote, respectively, Alice’s and Bob’s bit at that special coordinate. Let \(D\) be the sample-space duality that swaps the roles of Alice and Bob and reverses the coordinate order, let \(p^{\ast }\) be the dual protocol, and let \(z^{\ast }\) be the dual of an achievable \(Z\)-value \(z\). Then for every achievable \(Z\)-value \(z\) of \(p\) and every bit \(b\),
Equivalently, a sample \(\omega \) has \(Z_p(\omega ) = z\) with Bob’s special bit equal to \(b\) if and only if its dual \(D(\omega )\) has dual \(Z\)-value \(z^{\ast }\) with Alice’s special bit equal to \(b\).
Let \(n\) be a positive integer, and let \(p\) be a deterministic two-party communication protocol in which both Alice’s and Bob’s inputs are subsets of \(\{ 0,1,\dots ,n-1\} \) and the output is a single Boolean. Protocol duality produces from \(p\) a protocol over the same input and output types by exchanging the roles of Alice and Bob and reversing the coordinate order \(i \mapsto n-1-i\) in both players’ inputs, and it induces a map on syntactic transcripts that sends each execution transcript of \(p\) to the transcript of the dual protocol carrying the very same sequence of Boolean messages. This induced transcript map is injective.
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\). The dualization map sends a raw \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) value for \(p\) to the corresponding value for the dual protocol — the protocol obtained from \(p\) by exchanging the roles of Alice and Bob and reversing both coordinate sets — as follows: the transcript \(M\) is carried through the transcript-level map induced by this duality; the special coordinate \(T \in \{ 0, \dots , n-1\} \) is replaced by its reversal \(n - 1 - T\); the new “before” vector is the coordinate-reversal of \(Y_{\gt T}\); and the new “after” vector is the coordinate-reversal of \(X_{\lt T}\). This map is injective.
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) over disjointness inputs of length \(n\). An (achievable) \(Z\)-value of \(p\) is a tuple recording a transcript of \(p\), a special coordinate \(t \in \{ 0,\dots ,n-1\} \), and two Boolean vectors capturing Alice’s bits before \(t\) and Bob’s bits after \(t\), subject to the constraint that the tuple arises from some hard-distribution sample. Consider the dualization map that sends each such \(Z\)-value of \(p\) to the corresponding \(Z\)-value of the dual protocol, by dualizing the transcript (swapping the roles of Alice and Bob and reversing both coordinate sets), replacing the special coordinate \(t\) by its reverse \(n-1-t\), and interchanging the two Boolean vectors while reversing the coordinate order of each. This map is injective.
Fix \(n \ge 1\), let \(p\) be a deterministic two-party protocol for the disjointness problem on inputs of length \(n\), and let \(\mu \) be any measure on the hard sample space. Executing \(p\) on the input pair generated by a sample produces a syntactic transcript, and we regard this transcript as a random variable on the sample space. Then its Shannon entropy, computed with respect to \(\mu \), satisfies
where \(c(p)\) is the communication complexity of \(p\), i.e. the worst-case number of bits exchanged.
Fix a deterministic disjointness protocol \(p\) on inputs of length \(n\), an achievable value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), and a pair of bits \(b = (b_1, b_2)\). Write \(\mu \) for the uniform measure on the sample space of the hard distribution, and let \(F\) be the fiber of \(Z\) above \(z\), that is, the set of samples whose \(Z\)-value equals \(z\). Denote by \(X_T\) and \(Y_T\) Alice’s and Bob’s bits at the special coordinate \(T\). Then
Fix a positive integer \(n\) and a coordinate \(i \in \{ 0,1,\dots ,n-1\} \). Let \(R\) be the map that sends a pair \((a,b)\) — where \(a\) is a Boolean vector indexed by \(\{ 0,\dots ,i-1\} \) and \(b\) is a Boolean vector of length \(n\) — to
whose three components are all Boolean vectors of length \(n\); here \([\, \cdot \, ]\) is the Iverson bracket, so each component takes the value \(0\) at every index falling outside the range named in its own bracket. For every sample \(\omega \) of the hard distribution, let \(b\) be Bob’s full induced input vector of length \(n\) and let \(a\) be Alice’s induced input restricted to the coordinates \(\{ 0,\dots ,i-1\} \) (a Boolean vector on those indices, carrying no padding). Then
where \(x_{\lt i}\), \(y_{\lt i}\), and \(y_{\ge i}\) are the length-\(n\) Boolean vectors that equal Alice’s induced input at each index below \(i\), Bob’s induced input at each index below \(i\), and Bob’s induced input at each index at least \(i\), respectively, and are \(0\) at every other index.
Fix \(n \ge 1\) and an index \(i \in \{ 0, 1, \dots , n-1\} \), and consider Boolean strings valued in \(\mathbb {F}_2 = \{ 0,1\} \). Define a map that sends a pair \((x, y)\), where \(x \colon \{ 0,\dots ,i-1\} \to \mathbb {F}_2\) records the first \(i\) coordinates of Alice’s input and \(y \colon \{ 0,\dots ,n-1\} \to \mathbb {F}_2\) is Bob’s full input, to the triple \((u, v, w)\) of Boolean functions on all of \(\{ 0,\dots ,n-1\} \) given by \(u(j) = y(j)\) if \(j \lt i\) and \(u(j) = 0\) otherwise, \(v(j) = x(j)\) if \(j \lt i\) and \(v(j) = 0\) otherwise, and \(w(j) = y(j)\) if \(j \ge i\) and \(w(j) = 0\) otherwise; that is, \((x,y) \mapsto (y_{\lt i}, x_{\lt i}, y_{\ge i})\) with all out-of-range coordinates zero-padded. Then this map is injective.
Fix a positive integer \(n\) and a coordinate \(i \in \{ 0, 1, \dots , n-1\} \), and equip the sample space of the hard distribution with the measure obtained by conditioning on the event that the generated pair of input sets is disjoint. There are two ways to record Alice’s conditioning information at coordinate \(i\) from a sample \(\omega \): directly, as the pair \(\big(X_{\lt i}(\omega ),\, Y_{\ge i}(\omega )\big)\) formed from the sample’s own Alice-prefix and Bob-suffix bits; and indirectly, by first forming the generated disjoint coordinate-pair vector of \(\omega \) and then extracting the coordinate-vector conditioning \(\big(\)
Fix a positive integer \(n\), write the coordinates as \(i \in \{ 0,\dots ,n-1\} \), and consider the hard distribution on samples conditioned on the event that Alice’s input set and Bob’s input set are disjoint. For any coordinate \(i\), under this conditioned distribution the conditional mutual information between Alice’s bit \(X_i\) at coordinate \(i\) and Bob’s prefix \(Y_{\lt i}\) (his bits at all coordinates below \(i\)), given the pair \((X_{\lt i}, Y_{\ge i})\) formed by Alice’s prefix and Bob’s suffix, is zero:
Fix a positive integer \(n\), a deterministic communication protocol \(p\) for disjointness inputs on the coordinate set \(\{ 0,\dots ,n-1\} \), and a coordinate \(i\). Draw a hard sample and condition on the event that the induced input sets \(X\) and \(Y\) are disjoint; under this conditioned distribution let \(X_i\) denote Alice’s bit at coordinate \(i\), let \(M\) denote the transcript produced by executing \(p\) on the induced input pair, let \(X_{\lt i}\) denote Alice’s bits at the coordinates below \(i\), let \(Y_{\ge i}\) denote Bob’s bits at the coordinates at least \(i\), and let \(Y\) denote Bob’s full bit vector. Then the conditional mutual information between \(X_i\) and \(M\), given the partial pair \((X_{\lt i}, Y_{\ge i})\), is at most the conditional mutual information given the fuller pair \((X_{\lt i}, Y)\):
Fix a positive integer \(n\) and a coordinate \(i \in \{ 1,\dots ,n\} \). Under the hard distribution conditioned on the event that the generated input pair is disjoint, Alice’s bit at coordinate \(i\) agrees almost everywhere with the \(x\)-bit of the \(i\)-th entry of the disjoint coordinate-vector generated from the sample; that is, the two functions \(\omega \mapsto \)
Fix a coordinate index \(i\). Consider the hard distribution on samples \(\omega \) conditioned on the event that the generated input pair is disjoint. Then, almost surely with respect to this conditioned distribution, the two prefix functions on coordinates agree: the map sending each \(j\) to Bob’s bit \(\mathrm{yBit}(\omega , j)\) when \(j \lt i\) and to \(\mathrm{false}\) otherwise coincides with the map sending each \(j\) to Bob’s bit at coordinate \(j\) of the generated disjoint coordinate vector of \(\omega \) when \(j \lt i\) and to \(\mathrm{false}\) otherwise.
Let \(\omega \) be a sample of the hard distribution, recording a special coordinate \(T\), Alice’s and Bob’s bits \(x_T, y_T\) at that coordinate, and an assignment of the remaining coordinates. Toggling Alice’s special-coordinate bit twice — that is, applying to \(\omega \) the map that negates \(x_T\) while leaving \(T\), \(y_T\), and the other coordinates unchanged, and then applying that same map again — recovers \(\omega \) exactly.
Fix \(n \ge 1\), and let \(\phi \) be the map on hard samples that flips only Alice’s special-coordinate bit, sending a sample \(\omega = (T, x_T, y_T, \mathrm{other})\) to \((T, \lnot x_T, y_T, \mathrm{other})\). Then for every hard sample \(\omega \), the preimage of the singleton \(\{ \omega \} \) under \(\phi \) is exactly the singleton \(\{ \phi (\omega )\} \); that is, \(\phi ^{-1}(\{ \omega \} ) = \{ \phi (\omega )\} \).
Let \(n\) be a positive integer, and consider the set-disjointness function on subsets of \([n] = \{ 0, 1, \dots , n-1\} \), which sends a pair \(X, Y \subseteq [n]\) to \(\mathtt{true}\) precisely when \(X \cap Y = \emptyset \). Then its \(\tfrac {1}{32}\)-error public-coin randomized communication complexity \(R^{\mathrm{pub}}_{1/32}\) satisfies
where the left-hand side is the floor of the real number \(n / 2^{32}\) and the right-hand side is valued in \(\mathbb {N}_\infty \).
Fix a positive integer \(n\), a deterministic protocol \(p\) that computes set disjointness on pairs of subsets of \([n]\) (Alice and Bob each holding one subset), and a real number \(\gamma \). Draw a sample \(\omega \) from the hard distribution, and let \(Z(\omega )\) be the summary consisting of the transcript that \(p\) produces on the associated input pair \((X(\omega ),Y(\omega ))\), the special coordinate \(T\), Alice’s input bits at the coordinates below \(T\), and Bob’s input bits at the coordinates above \(T\). Call \(\omega \) good if, conditioned on the value of \(Z(\omega )\), the law of the special bit-pair \((x_T,y_T)\) lies within total variation distance \(2\gamma \) of the uniform distribution on \(\{ 0,1\} ^2\). Then
where \(\mathrm{err}(p)\) is the distributional error of \(p\) against the disjointness function under the input distribution induced by the hard distribution.
Fix a positive integer \(n\) and a deterministic two-party communication protocol \(p\) that attempts to decide set disjointness on pairs of subsets of \([n]=\{ 1,\dots ,n\} \), and let \(\gamma \in \mathbb {R}\). Draw a sample \(\omega \) from the hard distribution, which produces inputs \(X(\omega ),Y(\omega )\subseteq [n]\) together with a distinguished special coordinate \(T\), and let \(Z(\omega )\) be the summary statistic consisting of the transcript of \(p\) on \((X(\omega ),Y(\omega ))\), the coordinate \(T\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\). Call \(\omega \) good when the conditional law of the special bit-pair \((x_T,y_T)\), given the value of \(Z(\omega )\), lies within total variation distance \(2\gamma \) of the uniform law on \(\{ 0,1\} ^2\), and call \(\omega \) erroneous when \(p\) outputs a value different from the true disjointness of \(X(\omega )\) and \(Y(\omega )\). Then, with all probabilities taken under the hard distribution,
A hard sample over \(n \ge 1\) coordinates consists of a distinguished coordinate \(T \in \{ 0,1,\dots ,n-1\} \), two bits \(x_T, y_T \in \{ 0,1\} \), and an assignment giving each of the \(n\) coordinates one of the three tags *neither*, *left only*, or *right only*. The number of such samples is exactly \(n \cdot 4 \cdot 3^{n}\).
Fix a positive integer \(n\), and let \(\mathrm{HardSample}(n)\) be the finite sample space whose points \(\omega \) record a special coordinate \(\omega .T \in \{ 0,\dots ,n-1\} \), two bits \(\omega .x_T,\omega .y_T \in \{ 0,1\} \), and a coordinate vector \(\omega .\mathrm{other}\colon \{ 0,\dots ,n-1\} \to \mathrm{DisjointCoordinate}\). For any set \(\mathcal T\) of special coordinates, any sets \(\mathcal X,\mathcal Y\) of bits, and any set \(\mathcal O\) of coordinate vectors, the ambient measure of the rectangle event
taken as a real number, equals the product
where \(\mu _{\mathrm{Fin}}\) is the uniform law on the special coordinate, \(\mu _{\mathrm{bit}}\) is the uniform law on one bit, and \(\mu _{\mathrm{vec}}\) is the uniform law on coordinate vectors, each evaluated on the corresponding set as a real number.
Fix a positive integer \(n\), and consider the hard distribution on samples conditioned on the disjointness event — the event that the generated input pair \((X,Y)\) is disjoint. Under this conditioned distribution, the generated disjoint coordinate vector, viewed as a function \(\{ 1,\dots ,n\} \to \{ \text{neither}, \text{leftOnly}, \text{rightOnly}\} \), is identically distributed to a coordinate vector drawn from the uniform law on all such functions; that is, its law is exactly the uniform distribution over the \(3^n\) possible disjoint coordinate vectors.
Let \(n \ge 1\) and fix a coordinate \(i \in \{ 0, \dots , n-1\} \). Consider the hard distribution on samples, conditioned both on the event that the generated pair of sets is disjoint and on the special coordinate being equal to \(i\). Then the generated vector of disjoint coordinates, viewed as a map \(\{ 0,\dots ,n-1\} \to \{ \text{neither}, \text{leftOnly}, \text{rightOnly}\} \), has the same law as the identity map applied to a vector drawn uniformly at random from all such maps; that is, under this conditioning the generated vector is uniformly distributed over the \(3^n\) possible coordinate assignments.
Fix a positive integer \(n\) and a coordinate \(i \in \{ 0, 1, \dots , n-1\} \). Consider two ways of producing a triple. First, draw a hard-distribution sample conditioned on the event that Alice’s input set \(X\) and Bob’s input set \(Y\) are disjoint, and record the triple \(\bigl(x_i,\; Y_{\lt i},\; (X_{\lt i}, Y_{\ge i})\bigr)\), where \(x_i\) is Alice’s bit at coordinate \(i\), \(Y_{\lt i}\) is Bob’s bit vector on the coordinates \(j \lt i\) (set to \(\mathrm{false}\) elsewhere), \(X_{\lt i}\) is Alice’s bit vector on the coordinates \(j \lt i\), and \(Y_{\ge i}\) is Bob’s bit vector on the coordinates \(j \ge i\) (each padded with \(\mathrm{false}\) outside its range). Second, draw a vector of disjoint coordinates uniformly at random and form the same triple from Alice’s and Bob’s bits read off that vector. Then these two random triples are identically distributed.
Let \((\Omega , \mu )\) and \(\alpha \) be measurable spaces, let \(T \colon \Omega \to \Omega \) be a measure-preserving map (so that \(\mu (T^{-1}(A)) = \mu (A)\) for every measurable set \(A\)), and let \(F \colon \Omega \to \alpha \) be measurable. Then the random variables \(F\) and \(F \circ T\) are identically distributed under \(\mu \); that is, they induce the same pushforward measure on \(\alpha \).
Fix an integer \(n \ge 1\), and consider the hard distribution on samples conditioned on the event that the generated input pair \((X,Y)\) is disjoint. Under this conditional law, the pair consisting of the special coordinate \(T \in \{ 1,\dots ,n\} \) and the generated disjoint coordinate vector \(c \colon \{ 1,\dots ,n\} \to \{ \text{neither}, \text{leftOnly}, \text{rightOnly}\} \) has the same distribution as a pair of independent random variables in which \(T\) is uniform on \(\{ 1,\dots ,n\} \) and \(c\) is uniform on all \(3^{n}\) coordinate vectors. In particular \(T\) and \(c\) are independent, and each is uniformly distributed on its range.
Fix a positive integer \(n\), and consider pairs \((t, c)\) in which \(t\) is an element of \(\{ 0, 1, \dots , n-1\} \) and \(c\) is a disjoint-coordinate vector, that is, a function assigning to each of the \(n\) coordinates one of the three values "neither", "left only", or "right only". Draw \((t, c)\) from the product of the uniform law on the special coordinate and the uniform law on disjoint-coordinate vectors. Then the second component \(c\) of this pair is identically distributed to a disjoint-coordinate vector drawn directly from the uniform law on disjoint-coordinate vectors; equivalently, the second marginal of the product law coincides with that uniform law.
Fix a positive integer \(n\) and an index \(i \in \mathrm{Fin}\, n\), and let \(\mathrm{DisjointCoordinate} = \{ \text{neither}, \text{leftOnly}, \text{rightOnly}\} \) be the three-element type. On the product space \(\mathrm{Fin}\, n \times (\mathrm{Fin}\, n \to \mathrm{DisjointCoordinate})\) form the product of the uniform law on the special coordinate and the uniform law on disjoint-coordinate vectors, and condition it on the event that the first coordinate equals \(i\). Then under this conditional law the second coordinate is identically distributed to a uniformly random disjoint-coordinate vector.
Fix \(n \ge 1\), and let \(\omega \) be a sample of the hard distribution, from which one reads off Alice’s input set \(X(\omega )\) and Bob’s input set \(Y(\omega )\), packaged as the input pair \((X(\omega ), Y(\omega ))\). Form the dual sample by swapping the roles of Alice and Bob and reversing the coordinate order \(i \mapsto \bar\imath \). Then the input pair generated by the dual sample is obtained from \((X(\omega ), Y(\omega ))\) by exchanging the two sets and reversing the coordinate order of each: Alice receives the coordinate-reversal of \(Y(\omega )\) and Bob receives the coordinate-reversal of \(X(\omega )\), i.e.
Fix a positive integer \(n\), and let \(p\) be a deterministic two-party protocol whose parties each receive a subset of \(\{ 0, 1, \dots , n-1\} \) and which outputs a Boolean value. For any sample \(\omega \) of the hard distribution, write \(\mathrm{input}(\omega ) = (X(\omega ), Y(\omega ))\) for the pair of input sets it produces, and let \(t\) be the transcript obtained by executing \(p\) on \(\mathrm{input}(\omega )\). Then \(\mathrm{input}(\omega )\) belongs to the input set of \(t\); that is, the input pair generated by \(\omega \) lies among those input pairs whose execution of \(p\) follows exactly the path recorded by the transcript it produces.
Fix a positive integer \(n\) and a deterministic protocol \(p\) for disjointness inputs of length \(n\). Let \(z\) be an achievable value of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) for \(p\), and let \(\omega \) be a hard-distribution sample whose \(Z\)-variable equals \(z\). Then the input pair generated by \(\omega \) belongs to the input set of the transcript component of \(z\); that is, executing \(p\) on this input pair follows exactly the path recorded by that transcript.
Let \(\omega _X\) and \(\omega _Y\) be two hard-distribution samples that agree on their special coordinate \(T\), whose associated Alice-bit vectors before the special coordinate coincide, and whose associated Bob-bit vectors after the special coordinate coincide. Then the input pair generated by mixing Alice’s data from \(\omega _X\) with Bob’s data from \(\omega _Y\) is exactly \((X(\omega _X),\, Y(\omega _Y))\): the Alice set of the mixed sample equals Alice’s input set under \(\omega _X\), and the Bob set of the mixed sample equals Bob’s input set under \(\omega _Y\).
Fix a positive integer \(n\), and let \(S\) be a set of pairs \((A,B)\) of subsets of \(\{ 0,1,\dots ,n-1\} \). Then the real-valued probability that the pair of input sets \((X,Y)\) generated by the hard input distribution lies in \(S\) equals the real-valued measure, under the hard-distribution measure on the sample space, of the preimage \(\{ \omega : (X(\omega ),Y(\omega )) \in S\} \) — that is, the probability assigned to \(S\) by the hard input distribution is computed by pulling \(S\) back through the input map to the explicit hard sample space.
Fix an input length \(n \ge 1\), a deterministic communication protocol \(p\) for disjointness on inputs of length \(n\), and a real number \(\gamma \gt 0\). Draw a sample \(\omega \) from the hard distribution \(\mathbb {P}\); write \(M\) for the transcript produced by \(p\) on the associated input pair, write \(T\) for the special coordinate, and set \(Z = (M,\, T,\, X_{\lt T},\, Y_{\gt T})\), collecting the transcript together with \(T\), Alice’s bits strictly before \(T\), and Bob’s bits strictly after \(T\). For each value \(z\) of \(Z\), let \(\Delta (z)\) denote the total variation distance between the conditional law of Alice’s special bit \(X_T\) given \(Z = z\) under \(\mathbb {P}\) and the uniform law on one bit. Suppose the combined conditional mutual information
where both terms are computed under \(\mathbb {P}\) conditioned on the disjointness event \(D\) (that Alice’s and Bob’s input sets are disjoint). Then, with the expectation taken under \(\mathbb {P}\) conditioned on \(D\) together with \(Y_T = 0\),
Fix a positive integer \(n\) and a deterministic protocol \(p\) for the disjointness problem on inputs of length \(n\). Let \(\mu \) denote the hard distribution on samples, conditioned on the event that Alice’s set and Bob’s set are disjoint and on the event that Bob’s bit at the special coordinate \(T\) equals \(\mathtt{false}\); let \(\mathbb {E}_\mu \) denote expectation under \(\mu \), and let \(I_\mu \) denote conditional mutual information under \(\mu \). Let \(M\) denote the transcript produced by running \(p\) on the generated input pair, let \(X_T\) denote Alice’s bit at the special coordinate, and let \(Z = (M,\, T,\, X_{\lt T},\, Y_{\gt T})\) denote the fine variable consisting of the transcript together with the coarse data \((T, X_{\lt T}, Y_{\gt T})\). Then the \(\mu \)-average of the Kullback–Leibler divergence \(D(\cdot \, \| \, \cdot )\) from the conditional law of \(X_T\) given \(Z\) to the uniform distribution on \(\{ 0,1\} \) equals the conditional mutual information between \(X_T\) and \(M\) given the coarse data:
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), and work under the hard distribution conditioned on Alice’s and Bob’s generated input sets being disjoint and on Bob’s bit \(Y_T\) at the special coordinate \(T\) being \(\mathrm{false}\). Write \(M\) for the transcript that \(p\) produces on the generated input, let \(Z = (M,\, T,\, X_{\lt T},\, Y_{\gt T})\) be the associated fiber variable, and let \(U\) be the uniform law on one bit. Then the expectation, taken over this conditioned distribution, of the Kullback–Leibler divergence \(D_{\mathrm{KL}}\! \big(P_{X_T \mid Z}\, \big\| \, U\big)\) of the conditional law of Alice’s special bit \(X_T\) given \(Z\) relative to \(U\) equals the conditional mutual information \(I\big(X_T ; M \mid T,\, X_{\lt T},\, Y_{\ge T}\big)\) between \(X_T\) and \(M\) given Alice’s conditioning variable \((T,\, X_{\lt T},\, Y_{\ge T})\), computed under the same conditioned distribution.
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the \(n\)-coordinate disjointness problem, and let \(\mu \) be the hard distribution on samples \(\omega \) conditioned on the two generated input sets being disjoint and on Bob’s bit at the special coordinate \(T\) being false. Write \(X_T(\omega )\) for Alice’s bit at the special coordinate, and let \(Z(\omega )\) be the summary variable recording the transcript of \(p\) on the generated input pair together with the special coordinate \(T\), Alice’s coordinates strictly before \(T\), and Bob’s coordinates strictly after \(T\). For a value \(z\) of \(Z\), let \(\mathcal{L}(X_T \mid Z = z)\) denote the law of \(X_T\) under the hard distribution conditioned on the fiber \(\{ Z = z\} \), and let \(U\) be the uniform law on a single bit. Then
so that the \(\mu \)-average of the fiberwise Kullback–Leibler divergences from uniform equals the conditional Kullback–Leibler divergence of \(X_T\) given \(Z\), measured from uniform, under \(\mu \).
Fix a positive integer \(n\) and a deterministic protocol \(p\) for set disjointness on subsets of \(\{ 1,\dots ,n\} \). Let \(\mu \) denote the hard distribution over samples, conditioned both on the two generated input sets being disjoint and on Bob’s bit at the special coordinate being \(\mathrm{false}\). Let \(Z\) be the statistic sending each sample \(\omega \) to the tuple consisting of the transcript of \(p\) on its input pair, the special coordinate, Alice’s input bits strictly before that coordinate, and Bob’s input bits strictly after it; the map \(Z\) takes only finitely many values. For each such value \(z\), let \(\mathrm{KL}(z)\) be the real-valued Kullback–Leibler divergence, from the uniform distribution on one bit, of the conditional law of Alice’s special bit given \(Z = z\). Then the \(\mu \)-average of the KL cost, evaluated along \(Z\), splits as a finite fiber-weighted sum:
where \(z\) ranges over all values of \(Z\) and \(\mu \bigl(Z^{-1}(z)\bigr)\) is the \(\mu \)-measure of the fiber \(\{ \omega : Z(\omega ) = z\} \).
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) whose inputs are pairs of subsets of an \(n\)-element coordinate set. Let \(\mu \) be the hard distribution conditioned both on the two generated input sets being disjoint and on Bob’s bit at the special coordinate being \(\mathrm{false}\). For a sample \(\omega \), let \(Z(\omega )\) record the transcript of \(p\) on the generated input, the special coordinate \(T\), Alice’s input bits at the coordinates strictly before \(T\), and Bob’s input bits at the coordinates strictly after \(T\), and let \(\mathcal Z\) be the finite set of values it takes; for each \(z \in \mathcal Z\) write \(F_z = \{ \omega : Z(\omega ) = z\} \) for the corresponding fiber. For a value \(z\), let \(\kappa (z)\) be the Kullback–Leibler divergence, taken as a real number, of the law of Alice’s bit at the special coordinate — obtained from the ambient sample measure conditioned on \(F_z\) — from the uniform law on one bit. Then
where \(\mathcal L_\mu (\, \cdot \mid Z = z)\) denotes the law under \(\mu \) conditioned on the fiber \(F_z\).
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for set-disjointness, whose inputs are the pairs of subsets \(X, Y \subseteq \{ 0, 1, \dots , n-1\} \) produced by the hard distribution. Let \(\mu \) be the hard distribution conditioned both on the two generated sets being disjoint and on Bob’s bit \(Y_T\) at the special coordinate \(T\) being \(\mathrm{false}\), and for a sample \(\omega \) write \(\Pi (\omega )\) for the transcript of \(p\) run on the input pair \((X, Y)\) generated by \(\omega \). Let \(Z(\omega ) = \big(\Pi (\omega ),\, T,\, X_{\lt T},\, Y_{\gt T}\big)\) be the statistic recording the transcript, the special coordinate, Alice’s bits strictly before \(T\), and Bob’s bits strictly after \(T\). For each value \(z\) that \(Z\) can take, let \(D(z)\) be the Kullback–Leibler divergence, from the uniform distribution on one bit, of the conditional law of Alice’s special bit \(X_T\) given \(Z = z\), where the conditioning is taken with respect to the ambient uniform law on samples. Then
where the conditional mutual information on the right is computed under the hard distribution conditioned only on the two generated sets being disjoint.
Fix a positive integer \(n\) and a deterministic two-party communication protocol \(p\) for set-disjointness on \(n\) coordinates, and let \(\gamma \gt 0\). For a sample \(\omega \) of the hard distribution, write \(T\) for its special coordinate, \(X_T\) and \(Y_T\) for Alice’s and Bob’s bits at \(T\), \(\Pi \) for the transcript produced by running \(p\) on the induced input pair \((X(\omega ), Y(\omega ))\), and \(Z = (\Pi ,\, T,\, X_{\lt T},\, Y_{\gt T})\) for the transcript together with the special coordinate, Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\). Suppose that, under the hard distribution conditioned on the two input sets being disjoint, the total claimed information satisfies
Then, taking the expectation over the same disjoint-conditioned hard distribution further conditioned on \(Y_T\) being \(\mathrm{false}\), the mean Kullback–Leibler divergence of the conditional law of \(X_T\) given the observed value of \(Z\) from the uniform distribution on one bit satisfies
Fix a positive integer \(n\), and let \(\mathrm{DISJ}_n\) denote the set-disjointness function that sends a pair of subsets \(X, Y \subseteq [n]\) to \(\mathtt{true}\) precisely when \(X \cap Y = \emptyset \). Write \(\mathrm{R}^{\mathrm{pub}}_{1/32}(\mathrm{DISJ}_n) \in \mathbb {N}\cup \{ \infty \} \) for the public-coin randomized communication complexity of \(\mathrm{DISJ}_n\) at error \(\tfrac {1}{32}\), i.e. the least worst-case number of bits exchanged by any public-coin protocol that computes \(\mathrm{DISJ}_n\) correctly with probability at least \(\tfrac {31}{32}\) on every input. Then for every natural number \(k\) satisfying
one has \(k \lt \mathrm{R}^{\mathrm{pub}}_{1/32}(\mathrm{DISJ}_n)\).
Fix a positive integer \(n\), and consider the finite sample space of hard-distribution samples, each of which determines Alice’s input set \(X \subseteq \{ 1,\dots ,n\} \) and Bob’s input set \(Y \subseteq \{ 1,\dots ,n\} \). Under the ambient measure on this sample space, the disjoint event — the set of samples for which \(X\) and \(Y\) are disjoint — has nonzero measure.
Fix a positive integer \(n\), and draw a hard sample \(\omega \) uniformly at random from the finite sample space of the hard distribution on \(n\) coordinates. Then the disjointness event — the event that Alice’s input set \(X(\omega )\) and Bob’s input set \(Y(\omega )\) are disjoint — has probability exactly \(3/4\).
Fix an integer \(n \ge 1\), and equip the finite space of hard samples with its uniform probability measure: a hard sample records a special coordinate \(T \in \{ 1,\dots ,n\} \), two bits \(x_T, y_T\), and a disjoint-coordinate value at each of the \(n\) coordinates, so there are \(n\cdot 4\cdot 3^{\, n}\) of them, each carrying mass \(1/(n\cdot 4\cdot 3^{\, n})\). Then for every point \(z\) of the space \(\{ 1,\dots ,n\} \times \big(\{ 1,\dots ,n\} \to \text{DisjointCoordinate}\big)\times \text{DisjointCoordinate}\), the set of hard samples that lie in the disjointness event and whose disjoint model equals \(z\) has measure exactly
the mass of a single sample.
Fix \(n \ge 1\), and equip the finite sample space of hard samples with its uniform probability measure. Then for every subset \(S\) of this space, the measure of \(S\) equals the ratio
that is, the number of hard samples lying in \(S\) divided by the total number of hard samples.
Consider the finite sample space whose elements record a distinguished coordinate \(T\), two special bits \(x_T, y_T \in \{ 0,1\} \), and a labelling of every other coordinate by one of three values, equipped with the uniform probability measure. For any prescribed pair of bits \(b_x, b_y \in \{ 0,1\} \), the event that \(x_T = b_x\) and \(y_T = b_y\) has probability \(\tfrac {1}{4}\).
Fix a positive integer \(n\) and an index \(i \in \{ 0, 1, \dots , n-1\} \). Equip the finite sample space of hard samples with its uniform probability measure. Then the probability of the event that the special coordinate \(T\) equals \(i\) is
Fix \(n \ge 1\) and equip the sample space of the hard distribution with the uniform probability measure, so that a sample consists of a special coordinate \(T \in \{ 1,\dots ,n\} \), Alice’s and Bob’s bits \(x_T, y_T\) at that coordinate, and an independent uniformly chosen label in \(\{ \text{neither},\text{leftOnly},\text{rightOnly}\} \) at each of the \(n\) coordinates. For any fixed coordinate \(i \in \{ 1,\dots ,n\} \), the probability that the special coordinate satisfies \(T = i\) and that the generated sets \(X\) and \(Y\) are disjoint equals
Equip the finite space of hard samples with the uniform probability measure, where each sample consists of a special coordinate \(T \in \{ 1,\dots ,n\} \), two Boolean bits \(x_T\) and \(y_T\), and a value in \(\{ \text{neither},\text{leftOnly},\text{rightOnly}\} \) for every one of the \(n\) coordinates. Then for each fixed coordinate \(i \in \{ 1,\dots ,n\} \), the probability that the special coordinate equals \(i\) and that both special bits are true, that is \(T = i\) and \(x_T = y_T = \text{true}\), is exactly \(\dfrac {1}{4n}\).
Fix an integer \(n \ge 1\), and let the sample space consist of all tuples \((T, x_T, y_T, c)\), where \(T \in \{ 1,\dots ,n\} \) is a distinguished coordinate, \(x_T, y_T \in \{ \mathrm{true}, \mathrm{false}\} \) are two Boolean values, and \(c\) assigns to each coordinate one of the three labels *neither*, *leftOnly*, or *rightOnly*; equip this finite space with the uniform probability measure. Then the event that the distinguished coordinate is a genuine intersection, namely \(\{ x_T = \mathrm{true} \text{ and } y_T = \mathrm{true}\} \), has probability \(\tfrac 14\).
Consider the finite sample space of hard samples for a fixed number of coordinates \(n \ge 1\), equipped with the uniform probability distribution, where each sample records in particular Bob’s bit \(y_T \in \{ \mathrm{true}, \mathrm{false}\} \) at the special coordinate. Then the event that this bit takes the value \(\mathrm{false}\) has probability \(\tfrac {1}{2}\).
Consider the hard distribution’s sample space, whose points carry two special-coordinate bits \(x_T, y_T \in \{ 0,1\} \) (among other data), equipped with the uniform probability measure. The event that both special bits vanish, \(x_T = 0\) and \(y_T = 0\), has probability \(\tfrac {1}{4}\).
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the length-\(n\) set-disjointness problem, and let \(\gamma \gt 0\). Work under the hard distribution on samples \(\omega \), each of which determines a special coordinate \(T\), Alice’s and Bob’s input sets \(X\) and \(Y\) (with special bits \(x_T\) and \(y_T\)), and the transcript \(M\) produced by \(p\) on input \((X,Y)\). For a value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), let \(\delta _X(z)\) be the total variation distance between the conditional law of \(x_T\) given \(\{ Z = z\} \) and the uniform distribution on one bit. Suppose that, under the hard distribution conditioned on the disjointness event \(X \cap Y = \varnothing \),
Then the hard distribution assigns mass at most \(\gamma /2\) to the set of samples for which both special bits vanish, \(x_T = y_T = 0\), and \(\delta _X(Z) \gt \gamma \).
Fix a positive integer \(n\) and a deterministic two-party protocol \(p\) for the set-disjointness problem in which Alice and Bob hold subsets \(X, Y \subseteq \{ 1, \dots , n\} \). Draw a sample \(\omega \) uniformly from the finite hard-sample space; a draw fixes a special coordinate \(T\), Alice’s and Bob’s bits \(x_T\) and \(y_T\) at the special coordinate, and the two input sets \(X(\omega )\) and \(Y(\omega )\). Let \(M(\omega )\) be the transcript produced by \(p\) on the input \((X(\omega ), Y(\omega ))\), and let
record the transcript, the special coordinate, Alice’s bits at coordinates strictly below \(T\), and Bob’s bits at coordinates strictly above \(T\). For a value \(z\) taken by \(Z\), let \(\delta (z)\) be the total-variation distance between the uniform law on a single bit and the conditional law of \(y_T\) given \(Z = z\), the latter formed under the uniform law on the hard-sample space. Write \(I(\, \cdot \, ;\, \cdot \mid \cdot \, )\) for conditional mutual information computed under the measure obtained by conditioning the uniform law on the event that \(X(\omega )\) and \(Y(\omega )\) are disjoint. Let \(\gamma \gt 0\) be a real number for which
Then the probability, under the uniform law on the hard-sample space, that \(x_T = y_T = 0\) and \(\delta (Z(\omega )) \gt \gamma \) is at most \(\gamma /2\).
Fix \(n \ge 1\), and consider a sample from the hard distribution: a special coordinate \(T \in \{ 1,\dots ,n\} \), bits \(x_T, y_T \in \{ 0,1\} \) at that coordinate, and for every coordinate a label drawn from \(\{ \text{neither},\text{leftOnly},\text{rightOnly}\} \) that fixes the two players’ bits away from \(T\). This sample determines Alice’s set \(X = \{ \, i : \text{Alice's bit at } i = 1\, \} \) and Bob’s set \(Y = \{ \, i : \text{Bob's bit at } i = 1\, \} \). If Bob’s bit at the special coordinate satisfies \(y_T = 0\), then \(X\) and \(Y\) are disjoint.
Let \(p\) be a deterministic communication protocol for disjointness inputs of length \(n\), let \(\gamma \) be a real number, and let \(\omega \) be a sample of the hard distribution. Associated with \(\omega \) is a value \(z\) of the observation statistic, which records the transcript of \(p\) on the input pair induced by \(\omega \), the special coordinate, Alice’s input bits at the coordinates before it, and Bob’s input bits at the coordinates after it; condition the hard distribution on the fiber of samples producing this same \(z\). Suppose that, under this conditioning, the total variation distance between the law of Alice’s bit at the special coordinate and the uniform law on one bit is at most \(\gamma \), and likewise the total variation distance between the law of Bob’s bit at the special coordinate and the uniform law on one bit is at most \(\gamma \). Then the total variation distance between the conditional law of the special-coordinate bit-pair and the uniform law on a bit-pair is at most \(2\gamma \).
Let \(n\) be a positive integer, and let \(p\) be a deterministic communication protocol whose inputs are pairs \((A,B)\) of subsets of \(\{ 0,1,\dots ,n-1\} \). Then, for almost every sample \(\omega \) under the hard distribution conditioned on the event that its generated input pair is disjoint, the transcript that \(p\) produces when run on the input pair \((X(\omega ),Y(\omega ))\) generated directly by \(\omega \) equals the transcript that \(p\) produces when run on the input pair reconstructed from the disjoint coordinate vector generated by \(\omega \).
Fix a length \(n \ge 1\), let \(p\) be a deterministic communication protocol for disjointness inputs on the coordinate set \(\{ 0,\dots ,n-1\} \), and let \(\omega \) be a point of the hard sample space. Running the dual protocol of \(p\) — the protocol obtained by exchanging the roles of Alice and Bob and reversing both coordinate sets — on the dual sample of \(\omega \) — the sample obtained by swapping Alice’s and Bob’s bits and reversing coordinate order — produces exactly the transcript obtained from the transcript of \(p\) on the input pair generated by \(\omega \) by applying the transcript-level recoding induced by protocol duality (that is, swapping the two players’ nodes and then re-routing through the coordinate-reversal maps).
Fix a deterministic communication protocol \(p\) for disjointness on \(n\) coordinates, a value \(z\) of the statistic \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) recorded by the protocol (its transcript, the special coordinate, Alice’s input bits before the special coordinate, and Bob’s input bits after it), and two bits \(b_X, b_Y\). Let \(\omega _X\) and \(\omega _Y\) be hard-distribution samples that both lie in the fiber \(Z = z\), and suppose that Alice’s bit at the special coordinate of \(\omega _X\) equals \(b_X\) while Bob’s bit at the special coordinate of \(\omega _Y\) equals \(b_Y\). Then the mixed sample — taking Alice’s input from \(\omega _X\) and Bob’s input from \(\omega _Y\) — again lies in the fiber \(Z = z\), and its pair of special-coordinate bits is \((b_X, b_Y)\).
Fix a deterministic disjointness protocol \(p\) on inputs of length \(n\), and let \(z\) be an achievable value of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), whose fiber \(Z^{-1}(z)\) is the set of hard-distribution samples mapping to \(z\). Let \(\omega _{\mathrm{pair}}\) and \(\omega \) be two samples of this fiber, and suppose that the pair of bits at the special coordinate of \(\omega _{\mathrm{pair}}\), namely Alice’s bit and Bob’s bit there, equals \(b = (b_1, b_2)\). Form the mixed sample by taking Alice’s input from \(\omega _{\mathrm{pair}}\) and Bob’s input from \(\omega \) (keeping their common special coordinate \(T\), the bits of \(X\) before \(T\), and the bits of \(Y\) after \(T\)). Then this mixed sample again lies in \(Z^{-1}(z)\), and Alice’s bit at its special coordinate equals \(b_1\).
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the length-\(n\) disjointness problem, and consider the hard-distribution samples together with the derived variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) recording the transcript, the special coordinate, Alice’s bits before it, and Bob’s bits after it. Let \(z\) be an achievable value of \(Z\), let \(\omega \) and \(\omega '\) be two samples, and let \(b = (b_1, b_2)\) be a pair of Booleans. Suppose that both \(\omega \) and \(\omega '\) produce the value \(z\) (that is, each lies in the fiber \(Z^{-1}(z)\)), and that the special-coordinate bit-pair of \(\omega '\) equals \(b\), so that Alice’s and Bob’s bits at the special coordinate of \(\omega '\) are \(b_1\) and \(b_2\) respectively. Then the mixed sample formed by taking Alice’s data from \(\omega \) and Bob’s data from \(\omega '\) again produces the value \(z\), and its bit for Bob at the special coordinate equals \(b_2\).
For two samples \(\omega ,\omega '\) of the hard distribution, let \(\omega \ast \omega '\) denote their *mixture*: the sample built by taking Alice’s input from \(\omega \) and Bob’s input from \(\omega '\). Let \(\omega _X\) and \(\omega _Y\) be two samples that share the same special coordinate \(T\), whose Alice-bits before the special coordinate agree, and whose Bob-bits after the special coordinate agree. Then
Fix a deterministic protocol \(p\) for disjointness inputs of length \(n\) and a value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), recording the transcript, the special coordinate, Alice’s input bits before it and Bob’s input bits after it. Let \(\omega _X\) and \(\omega _Y\) be two hard-distribution samples that both lie in the fiber above \(z\), so that each realizes \(Z = z\), and suppose moreover that Alice’s bit at the special coordinate of \(\omega _X\) equals \(b_X\) and Bob’s bit at the special coordinate of \(\omega _Y\) equals \(b_Y\). Then the mixed sample obtained by taking Alice’s coordinates from \(\omega _Y\) together with Bob’s coordinates from \(\omega _X\) again lies in the fiber above \(z\).
Fix \(n \ge 1\) and a deterministic disjointness protocol \(p\) on the coordinate set \(\{ 0,\dots ,n-1\} \). For a hard-distribution sample \(\omega \), let \(X(\omega ), Y(\omega ) \subseteq \{ 0,\dots ,n-1\} \) denote Alice’s and Bob’s input sets, and recall that the \(Z\)-variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) records, among other data, the transcript \(M\) that \(p\) produces on the input pair \((X(\omega ), Y(\omega ))\). Suppose two hard-distribution samples \(\omega \) and \(\omega '\) yield the same value \(z\) of this \(Z\)-variable under \(p\), and let \(M\) be the transcript component of \(z\). Then both mixed input pairs \((X(\omega '), Y(\omega ))\) and \((X(\omega ), Y(\omega '))\) belong to the input set of \(M\), that is, the set of all input pairs whose execution of \(p\) follows the transcript \(M\).
Fix a positive integer \(n\) and take the hard distribution on samples conditioned on the event that Alice’s set \(X\) and Bob’s set \(Y\) are disjoint, and further conditioned on Bob’s special bit \(Y_T\) being false. Under this conditional distribution the mutual information between Alice’s special bit \(X_T\) and the coarse variable \((T,\, X_{\lt T},\, Y_{\gt T})\) is zero:
so that \(X_T\) is independent of \((T,\, X_{\lt T},\, Y_{\gt T})\).
Let \(n\) be a positive integer and let \(p\) be a deterministic two-party communication protocol for disjointness inputs of length \(n\). Let \(\mu \) denote the hard distribution on inputs, conditioned both on the event that Alice’s and Bob’s generated input sets are disjoint and on Bob’s bit at the special coordinate being \(\mathrm{false}\). Write \(T\) for the special coordinate, \(X_T\) for Alice’s bit at coordinate \(T\), \(X_{\lt T}\) for the string of Alice’s input bits at coordinates below \(T\) (set to \(\mathrm{false}\) at all other coordinates), \(Y_{\gt T}\) for the string of Bob’s input bits at coordinates above \(T\) (set to \(\mathrm{false}\) at all other coordinates), and \(M\) for the transcript that \(p\) produces on the sampled input pair. Then, with every information quantity computed under \(\mu \),
Let \(\omega \) be a sample of the hard distribution on \(n\) coordinates, with special coordinate \(T\), and let \(i\) be any coordinate with \(i \neq T\). Then Alice’s bit and Bob’s bit at coordinate \(i\) are never both equal to \(\mathrm{true}\); that is, it is not the case that both the Alice-bit at \(i\) and the Bob-bit at \(i\) hold.
Let \(\omega _X\) and \(\omega _Y\) be two samples of the hard distribution that share the same special coordinate \(T\), whose Alice-before-special vectors agree and whose Bob-after-special vectors agree. Then for every coordinate \(i \neq T\), it is never the case that Alice’s bit at \(i\) under \(\omega _X\) and Bob’s bit at \(i\) under \(\omega _Y\) are both true.
Fix a positive integer \(n\), and let \(p\) be a deterministic two-party communication protocol whose inputs are subsets \(X, Y \subseteq [n]\) (held by Alice and Bob, respectively) and whose output is Boolean. Consider the hard input distribution generated as follows: choose a special coordinate \(T \in [n]\), bits \(x_T, y_T \in \{ 0,1\} \), and, for every coordinate \(i \ne T\), a label in \(\{ \text{neither}, \text{left only}, \text{right only}\} \), all uniformly and independently; a "left only" (resp. "right only") label places \(i\) in \(X\) (resp. \(Y\)) and in no other set, "neither" places it in neither, while at \(T\) the bits \(x_T, y_T\) decide membership of \(T\) in \(X\) and in \(Y\). Suppose that against this distribution the protocol \(p\) computes set-disjointness — outputting whether \(X \cap Y = \emptyset \) — with distributional error at most \(1/32\), i.e. the probability that the output of \(p\) on \((X, Y)\) disagrees with the disjointness indicator is at most \(1/32\). Then, writing \(\Pi \) for the transcript produced by running \(p\) on \((X, Y)\) and conditioning throughout on the event that \(X\) and \(Y\) are disjoint,
where \(X_{\lt T}\) (resp. \(X_{\le T}\)) records Alice’s membership bits at the coordinates strictly below \(T\) (resp. at most \(T\)), \(Y_{\ge T}\) (resp. \(Y_{\gt T}\)) records Bob’s membership bits at the coordinates at least \(T\) (resp. strictly above \(T\)), and all remaining coordinates are set to \(0\).
Fix a positive integer \(n\) and a deterministic two-party protocol \(p\) for the set-disjointness problem on subsets of \(\{ 1,\dots ,n\} \). The hard distribution is the uniform probability measure on the finite sample space; a sample \(\omega \) determines a special coordinate \(T \in \{ 1,\dots ,n\} \), Alice’s input set \(X\) and Bob’s input set \(Y\), and the two special-coordinate bits \(X_T\) and \(Y_T\). Let \(M\) denote the transcript that \(p\) produces on the pair \((X,Y)\), and let \(Z = (M,\, T,\, X_{\lt T},\, Y_{\gt T})\) be the statistic recording the transcript, the special coordinate, Alice’s input bits strictly before the special coordinate, and Bob’s input bits strictly after the special coordinate. Say that a value \(z\) is good when the law of \((X_T,Y_T)\) under the hard distribution conditioned on the fiber \(\{ Z = z\} \) lies within total variation distance \(2\gamma \) of the uniform distribution on \(\{ 0,1\} ^2\). Suppose \(\gamma \gt 0\) and that, under the hard distribution conditioned on the event that \(X\) and \(Y\) are disjoint,
Then the hard-distribution probability that \(Z\) is good is at least \(\tfrac 14\, (1-4\gamma )\).
Fix a deterministic disjointness protocol \(p\) on inputs of length \(n\), a real number \(\gamma \), and a value \(z\) of the transcript-statistic. Suppose \(z\) is *good* for \(\gamma \), meaning that the conditional special-pair law given \(z\) — the distribution of the special-coordinate bit-pair \((x_T, y_T)\) conditioned on the value \(z\) — lies within total variation distance \(2\gamma \) of the uniform law on bit-pairs. Then for every pair \(b \in \{ 0,1\} \times \{ 0,1\} \), this conditional special-pair law assigns to the singleton \(\{ b\} \) a mass of at least \(\tfrac 14 - 2\gamma \).
Fix a positive integer \(n\) and consider the set-disjointness problem on subsets of \([n]\) under the hard distribution: a sample \(\omega \) specifies a special coordinate \(T \in [n]\), a bit-pair \((x_T, y_T)\) at \(T\), and at each remaining coordinate one of the three options *neither*, *Alice only*, or *Bob only*; from \(\omega \) one reads off Alice’s set \(X(\omega )\) and Bob’s set \(Y(\omega )\). Let \(p\) be a deterministic communication protocol with Boolean output on such inputs, and let \(Z(\omega )\) be the statistic recording the transcript of \(p\) on \((X(\omega ), Y(\omega ))\), the special coordinate \(T\), Alice’s bits at the coordinates below \(T\), and Bob’s bits at the coordinates above \(T\). Let \(\gamma \in \mathbb {R}\), and suppose the value \(z\) of \(Z\) is *good*, meaning that the total variation distance between the conditional law of the special bit-pair \((x_T, y_T)\) given \(Z = z\) and the uniform distribution on \(\{ 0,1\} ^2\) is at most \(2\gamma \). Then, conditioning on the event \(Z = z\), the probability that \(p\) errs on disjointness — that its output on \((X(\omega ), Y(\omega ))\) disagrees with whether \(X(\omega )\) and \(Y(\omega )\) are disjoint — is at least \(\tfrac {1}{4} - 2\gamma \).
Fix a positive integer \(n\) and a deterministic disjointness protocol \(p\) on inputs of length \(n\), and let \(z\) and \(z'\) be two raw \(Z\)-values for \(p\) — each consisting of a syntactic transcript of \(p\), a special coordinate in \(\{ 0,1,\dots ,n-1\} \), and two Boolean vectors indexed by \(\{ 0,1,\dots ,n-1\} \) recording the bits before and after that coordinate. If \(z\) and \(z'\) have the same transcript, the same special coordinate, the same "before" vector, and the same "after" vector, then \(z = z'\).
Fix an integer \(n \ge 1\), let \(p\) be a deterministic communication protocol for disjointness inputs of length \(n\), and let \(\omega \) be a sample of the hard distribution. Then forming the raw \(Z\)-variable of the dual protocol at the dual sample yields exactly the dual of the raw \(Z\)-variable of \(p\) at \(\omega \); that is,
where \(Z(p,\omega )\) is the raw \(Z\)-variable — recording the transcript of \(p\) on the input pair generated by \(\omega \), the special coordinate of \(\omega \), Alice’s bits strictly before the special coordinate, and Bob’s bits strictly after it — and where \(\mathrm{dual}(\cdot )\) denotes, respectively, the dual protocol, the dual hard sample, and the dual raw \(Z\)-value, each obtained by swapping the roles of Alice and Bob and reversing the coordinate order.
For a positive integer \(n\), let \(\rho \colon \{ 0,1\} ^n \to \{ 0,1\} ^n\) be the map that reverses the coordinate order of a Boolean vector, sending \(v = (v_0, v_1, \dots , v_{n-1})\) to \((v_{n-1}, \dots , v_1, v_0)\). Then \(\rho \) is injective: whenever two Boolean vectors have the same reversal, they are equal.
For any boolean vector \(v = (v_0, v_1, \ldots , v_{n-1})\) of length \(n \ge 1\), reversing its coordinate order twice returns the original vector: if \(v'\) denotes the vector obtained by reversing the coordinate order of \(v\), then reversing the coordinate order of \(v'\) recovers \(v\) exactly.
Fix a positive integer \(n\), and for a subset \(S \subseteq \{ 0, 1, \dots , n-1\} \) let \(\overline{S} = \{ \, i : (n-1) - i \in S \, \} \) be the subset obtained by reversing the coordinate order. Then applying this reversal twice returns the original set: \(\overline{\overline{S}} = S\) for every \(S \subseteq \{ 0, 1, \dots , n-1\} \).
Fix \(n \ge 1\) and a deterministic protocol \(p\) for disjointness inputs of length \(n\). Let \(\omega \) be a sample of the hard distribution, and let \(z\) be an achievable value of the \(Z\)-variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) for \(p\). If \(z\) is exactly the value of the \(Z\)-variable at \(\omega \), then running \(p\) on Alice’s input set \(X(\omega )\) and Bob’s input set \(Y(\omega )\) produces the same Boolean output as the output determined by \(z\) through its transcript leaf:
Let \(n \ge 1\), and let \(b_x, b_y \in \{ \mathrm{false}, \mathrm{true}\} \) be prescribed bits that are not both \(\mathrm{true}\). Then every hard-distribution sample whose special coordinate carries Alice’s bit \(b_x\) and Bob’s bit \(b_y\) generates a disjoint input pair: the special-bits event for \((b_x, b_y)\) is contained in the event that Alice’s set \(X\) and Bob’s set \(Y\) are disjoint.
Fix a deterministic protocol \(p\) for disjointness inputs of length \(n\), and recall that the \(Z\)-variable of a hard-distribution sample \(\omega \) records the tuple \((M, T, X_{\lt T}, Y_{\gt T})\) consisting of the protocol’s transcript on the input generated by \(\omega \), the special coordinate \(T\), and the padded bit-vectors of Alice’s input before \(T\) and Bob’s input after \(T\). If two hard-distribution samples \(\omega \) and \(\omega '\) yield the same value \(z\) of this \(Z\)-variable, then they select the same special coordinate; that is, \(\omega .T = \omega '.T\).
Fix a disjointness protocol \(p\) on inputs of length \(n\), and recall that the special coordinate of a hard-distribution sample \(\omega \) is the index \(T\) selected by that sample. Let \(z\) be an achievable value of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), whose special-coordinate component records the index \(T\). If the \(Z\) value produced by \(\omega \) under \(p\) equals \(z\), then the special coordinate of \(\omega \) coincides with the special-coordinate component of \(z\).
Fix a positive integer \(n\) and a coordinate \(i \in \{ 0,1,\dots ,n-1\} \). The set of hard samples whose special coordinate equals \(i\) — that is, the preimage under the special-coordinate map of the singleton \(\{ i\} \) — coincides exactly with the event that the special coordinate takes the value \(i\).
Fix a positive integer \(n\), and let \(\omega \) be a point of the hard sample space on \(n\) coordinates. Passing to the dualized sample — which swaps the roles of Alice and Bob and reverses the coordinate order — turns Alice’s bit at the special coordinate into Bob’s bit at the special coordinate of \(\omega \): Alice’s special bit of the dual of \(\omega \) equals Bob’s special bit of \(\omega \).
Let \(\omega \) be a sample from the hard distribution, and write \(X_T(\omega )\) for Alice’s bit at the special coordinate of \(\omega \). Let \(\omega '\) be the sample obtained from \(\omega \) by toggling Alice’s special-coordinate bit while leaving all other data unchanged. Then \(X_T(\omega ') = \lnot \, X_T(\omega )\); that is, reading Alice’s special bit after the flip yields the negation of Alice’s special bit before it.
Fix a positive integer \(n\), and let \(\omega _X\) and \(\omega _Y\) be two samples from the hard distribution, each carrying a special coordinate together with Alice’s and Bob’s bits there and a disjoint coordinate on every other index. Form the mixed sample that takes Alice’s data from \(\omega _X\) and Bob’s data from \(\omega _Y\). Then Alice’s bit at the special coordinate of the mixed sample equals Alice’s bit at the special coordinate of \(\omega _X\).
Let \(\omega \) be a sample of the hard distribution, and let \(\omega '\) be its dual, obtained by exchanging the roles of Alice and Bob and reversing the order of the coordinates. Then Bob’s bit at the special coordinate of \(\omega '\) equals Alice’s bit at the special coordinate of \(\omega \).
Fix a positive integer \(n\) and work over the hard-distribution sample space, whose samples \(\omega \) record a special coordinate \(T \in \{ 0,\dots ,n-1\} \) together with enough data to determine Alice’s and Bob’s bits at every coordinate. Let \(Y_T(\omega )\) denote Bob’s bit at the special coordinate, and let the conditioning variable be the triple \(\bigl(T,\, X_{\lt T},\, Y_{\ge T}\bigr)\) consisting of the special coordinate, Alice’s bits strictly before it (padded with \(\mathrm{false}\) elsewhere), and Bob’s bits from it onward (padded with \(\mathrm{false}\) elsewhere). Then the set of samples with \(Y_T = \mathrm{false}\) coincides exactly with the set of samples whose conditioning variable \((t,x,y)\) satisfies \(y(t) = \mathrm{false}\).
Let \(\omega \) be a sample of the hard distribution, and let \(\omega '\) be obtained from \(\omega \) by flipping only Alice’s special-coordinate bit (toggling \(X_T\) while leaving \(T\), the coordinate types, and Bob’s bits unchanged). Then Bob’s bit at the special coordinate is the same for both samples: the special-\(Y\) value of \(\omega '\) equals the special-\(Y\) value of \(\omega \).
Let \(\omega _X\) and \(\omega _Y\) be two samples of the hard distribution, and form the mixed sample that takes Alice’s input from \(\omega _X\) and Bob’s input from \(\omega _Y\). Then Bob’s bit at the special coordinate of the mixed sample equals Bob’s bit at the special coordinate of \(\omega _Y\).
Fix \(n \ge 1\), and let a hard sample \(\omega \) record a special coordinate \(T \in \{ 0,\dots ,n-1\} \), Alice’s and Bob’s bits \(x_T, y_T \in \{ \mathrm{false},\mathrm{true}\} \) at that coordinate, and a labelling of every coordinate by one of the three disjointness states. Let its dual be the hard sample obtained by reversing the coordinate order (sending each index \(i\) to \(n-1-i\)), swapping the sides of every coordinate label, and interchanging Alice’s and Bob’s special bits, so that the dual’s Alice-bit is \(y_T\) and its Bob-bit is \(x_T\). Then the dual of \(\omega \) has both of its special bits equal to \(\mathrm{false}\) if and only if \(\omega \) itself has both special bits equal to \(\mathrm{false}\).
Fix \(n \ge 1\) and consider a hard-distribution sample \(\omega \): a distinguished coordinate \(T \in \{ 1,\dots ,n\} \), two special bits \(x_T, y_T \in \{ 0,1\} \), and, at every coordinate, a disjoint coordinate-type (neither, left-only, or right-only) that determines Alice’s and Bob’s bits there. If both special bits vanish, \(x_T = 0\) and \(y_T = 0\), then Alice’s input set \(X(\omega )\) and Bob’s input set \(Y(\omega )\) are disjoint. Equivalently, the event that both special-coordinate bits are zero is contained in the event that the generated input pair is disjoint.
Fix a positive integer \(n\) and a deterministic protocol \(p\) for length-\(n\) disjointness inputs, and work throughout under the hard distribution on samples conditioned on the event that the generated input pair is disjoint. Under this distribution write \(X = (X_1, \ldots , X_n)\) for Alice’s input vector, \(Y\) for Bob’s input vector, \(M\) for the transcript produced by running \(p\) on the input pair \((X, Y)\), and \(X_{\lt i}\) for the strict prefix \((X_1, \ldots , X_{i-1})\) of Alice’s vector. Then
where each \(I(\, \cdot : \cdot \mid \cdot \, )\) denotes conditional mutual information taken with respect to the conditioned distribution.
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the disjointness problem on inputs of length \(n\). Let \(\mu \) be the hard distribution on samples conditioned on the event that Alice’s input set and Bob’s input set are disjoint, and, for a sample governed by \(\mu \), write \(M\) for the transcript produced by running \(p\) on the sample’s input pair, write \(X_i\) for Alice’s bit at coordinate \(i\), write \(X_{\lt i}\) for Alice’s bits at all coordinates below \(i\), write \(Y_{\ge i}\) for Bob’s bits at the coordinates from \(i\) onward, and write \(Y\) for Bob’s full bit vector. With every mutual information computed under \(\mu \),
Fix a positive integer \(n\) and a deterministic protocol \(p\) for the \(n\)-coordinate disjointness problem. Draw an input from the hard distribution conditioned on the event that Alice’s and Bob’s sets are disjoint, and write \(X = (X_1,\dots ,X_n)\) and \(Y = (Y_1,\dots ,Y_n)\) for Alice’s and Bob’s Boolean input vectors and \(M\) for the transcript produced by running \(p\) on that input pair; for each coordinate \(i\) let \(X_i\) be Alice’s bit there, \(X_{\lt i}\) her vector of bits at earlier coordinates, and \(Y_{\ge i}\) Bob’s vector of bits at \(i\) and later coordinates. Then, with every mutual information taken under this disjointness-conditioned distribution,
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the disjointness problem on inputs that are subsets of an \(n\)-element set. For a hard sample \(\omega \), let \(Z(\omega )\) be the record consisting of the transcript produced by running \(p\) on the induced input pair, the special coordinate \(T\), Alice’s input bits at the coordinates strictly below \(T\) (padded by \(\mathrm{false}\) elsewhere), and Bob’s input bits at the coordinates strictly above \(T\). Write \(\nu _{Z(\omega )}\) for the conditional law of Alice’s special bit \(X_T\) given the value of \(Z(\omega )\) — the pushforward, under the ambient uniform law on hard samples restricted to the fibre \(\{ Z = Z(\omega )\} \), of the map \(\omega \mapsto X_T\) — and let \(U\) denote the uniform distribution on a single bit. Let \(\mu \) be the hard distribution conditioned both on the induced input pair being disjoint and on Bob’s bit at the special coordinate being \(\mathrm{false}\). Then
where \(\mathrm{TV}\) is the total variation distance and \(D_{\mathrm{KL}}\) the Kullback–Leibler divergence.
Let \(n\) be a positive integer, let \(p\) be a deterministic communication protocol for the disjointness problem on inputs of length \(n\), and let \(z\) be an achievable value of the auxiliary variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) for \(p\). Write \(\mu _z\) for the conditional law of Alice’s bit at the special coordinate given \(Z = z\), and let \(\upsilon \) denote the uniform distribution on a single bit. Then, taking the total variation distance \(\mathrm{TV}(\mu _z, \upsilon )\) and the Kullback–Leibler divergence \(D_{\mathrm{KL}}(\mu _z \, \| \, \upsilon )\) (as an extended real, read off as a real number), the two satisfy
Fix \(n \ge 1\) and a deterministic two-party communication protocol \(p\) for disjointness inputs of length \(n\), and let \(z\) be an achievable value of the transcript-and-side-information variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\). Write \(\mu _z\) for the conditional law of Alice’s bit \(X_T\) at the special coordinate given that \(Z = z\) — that is, the pushforward under the map \(\omega \mapsto X_T(\omega )\) of the uniform measure on the sample space conditioned on the fiber \(\{ Z = z\} \) — and write \(\nu \) for the uniform law on a single bit. Then twice the squared total variation distance between \(\mu _z\) and \(\nu \) is at most their Kullback–Leibler divergence,
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\). Consider the hard distribution over samples conditioned on the generated input pair \((X, Y)\) being disjoint, and let \(M\) be the transcript that \(p\) produces on that input, \(X_T\) Alice’s bit at the special coordinate \(T\), and \((T, X_{\lt T}, Y_{\ge T})\) Alice’s corrected special-coordinate conditioning variable. Writing \(I_D(X_T : M \mid T, X_{\lt T}, Y_{\ge T})\) for the conditional mutual information under this disjoint-conditioned distribution, and \(I_{D_0}(X_T : M \mid T, X_{\lt T}, Y_{\ge T})\) for the same conditional mutual information under the distribution obtained by further conditioning on Bob’s special bit \(Y_T\) being false, one has
Let \(\mu \) denote the uniform probability law on the two-element type \(\mathrm{Bool}\). Then for every subset \(S \subseteq \mathrm{Bool}\), the real-valued measure of \(S\) equals \(|S|/2\), where \(|S|\) is the number of Boolean values lying in \(S\).
Under the uniform probability law on the two Boolean values, each individual outcome \(b\) carries mass exactly \(\tfrac {1}{2}\); that is, the singleton event \(\{ b\} \) has probability \(\tfrac {1}{2}\) for either choice of \(b\).
Write the uniform probability distribution on a single bit \(\{ 0,1\} \), which assigns mass \(\tfrac 12\) to each value. Then the uniform probability distribution on the four ordered bit-pairs \(\{ 0,1\} \times \{ 0,1\} \), which assigns mass \(\tfrac 14\) to each pair, is equal to the product of two independent copies of the single-bit uniform distribution.
Let a *disjoint coordinate* be one of the three symbols \(\{ \text{neither}, \text{leftOnly},\ \text{rightOnly}\} \), and let \(\mu \) be the uniform law on this three-element set. Then for every disjoint coordinate \(c\), the measure of the singleton \(\{ c\} \) satisfies \(\mu (\{ c\} ) = \tfrac {1}{3}\).
Fix \(n \ge 1\) and draw a random vector \(C = (C_1, \dots , C_n)\) whose \(n\) coordinates are independent, each uniform over the three states of a disjoint coordinate-pair: one in which both parties’ bits are \(0\), one carrying only Alice’s bit, and one carrying only Bob’s bit. For each coordinate \(k\) let \(X_k\) be Alice’s bit, equal to \(1\) exactly in the second state, and let \(Y_k\) be Bob’s bit, equal to \(1\) exactly in the third; in particular \(X_k\) and \(Y_k\) are never simultaneously \(1\). Fix a coordinate \(i\) and form Alice’s conditioning vector \(A = (A_1, \dots , A_n)\) by setting \(A_k = X_k\) for \(k \lt i\) and \(A_k = Y_k\) for \(k \ge i\). Then the conditional mutual information between \(X_i\) and the tuple \((Y_k)_{k \lt i}\) of Bob’s bits at the coordinates strictly before \(i\), given \(A\), vanishes:
Fix a positive integer \(n\) and draw \(C = (C_1,\dots ,C_n)\) uniformly at random from \(\{ \text{neither},\text{leftOnly},\text{rightOnly}\} ^n\), so that the \(n\) coordinates are independent and each is uniform on the three possibilities. For each index \(j\) set Alice’s bit \(X_j = \mathbf{1}[C_j = \text{leftOnly}]\) and Bob’s bit \(Y_j = \mathbf{1}[C_j = \text{rightOnly}]\). Then for every coordinate \(i\), the conditional mutual information, taken under this uniform law, between Alice’s bit at \(i\) and Bob’s bits strictly before \(i\), given the pair consisting of Alice’s bits strictly before \(i\) and Bob’s bits from \(i\) onward, vanishes:
Fix a positive integer \(n\), and consider length-\(n\) vectors whose entries are disjoint coordinates, each drawn from the three-element type \(\{ \text{neither}, \text{leftOnly}, \text{rightOnly}\} \). Then the uniform probability law on the set of all such vectors coincides with the product measure of \(n\) independent copies of the uniform law on a single disjoint coordinate.
Let \(n\) be a positive integer, and equip the finite set of functions \(x\colon \{ 0,1,\dots ,n-1\} \to \{ \text{neither},\text{leftOnly},\text{rightOnly}\} \) with the uniform probability law (each of the \(3^n\) such functions is equally likely). Then the coordinate maps \(x \mapsto x(i)\), for \(i = 0,1,\dots ,n-1\), are mutually independent random variables.
Fix a positive integer \(n\), a coordinate \(i \in \{ 0,1,\dots ,n-1\} \), and a bit vector \(b \colon \{ 0,\dots ,n-1\} \to \{ 0,1\} \). Draw a coordinate vector uniformly from the \(n\)-fold product of the three-element set whose members mark a coordinate as belonging to neither party, to Alice only, or to Bob only; write \(X_k = 1\) exactly when coordinate \(k\) belongs to Alice only and \(Y_k = 1\) exactly when it belongs to Bob only. Condition on the event that Alice’s conditioning bit vector — whose \(k\)-th entry is \(X_k\) when \(k \lt i\) and \(Y_k\) when \(k \ge i\) — equals \(b\). Under this conditional law, Alice’s bit \(X_i\) at coordinate \(i\) and the vector of Bob’s earlier bits, equal to \(Y_j\) for each \(j \lt i\) and to \(0\) at every \(j \ge i\), are independent random variables.
Fix a positive integer \(n\), and draw a random vector of \(n\) disjoint-coordinate symbols by sampling each of its \(n\) coordinates independently and uniformly from the three values *neither*, *left-only*, and *right-only*. For a coordinate index \(k\) write \(X_k\) for Alice’s bit at that coordinate (equal to \(1\) exactly when the \(k\)-th symbol is *left-only*) and \(Y_k\) for Bob’s bit (equal to \(1\) exactly when it is *right-only*). Then for every pair of indices \(i,k \in \{ 0,1,\dots ,n-1\} \) and every bit \(b \in \{ 0,1\} \), the event that Alice’s conditioning bit at coordinate \(k\) under cut-off \(i\) — namely \(X_k\) when \(k\lt i\) and \(Y_k\) otherwise — takes the value \(b\) has strictly positive probability under this uniform law.
Let \(n \ge 1\), and consider the finite set of vectors \(x \colon \{ 1,\dots ,n\} \to \{ \text{neither}, \text{left only}, \text{right only}\} \) assigning to each coordinate one of three labels, of which there are \(3^n\) in total. Under the uniform probability law on this set, every subset \(S\) of such vectors has measure
where \(\left\lvert S\right\rvert \) denotes the number of vectors lying in \(S\).
Fix a positive integer \(n\) and an index \(i \in \{ 0,\dots ,n-1\} \), and let the coordinate vector \(c = (c_0,\dots ,c_{n-1})\) be drawn from the uniform law on \(\{ \text{neither},\text{leftOnly},\text{rightOnly}\} ^{n}\). For each coordinate \(k\) form the pair \(\bigl(R_i(k,c),\, A_i(k,c)\bigr)\), where \(A_i(k,c)\) is Alice’s conditioning bit — the \(X\)-bit of \(c_k\) when \(k\lt i\) and the \(Y\)-bit of \(c_k\) when \(k\ge i\) — and \(R_i(k,c)\) is the corresponding residual bit — the \(Y\)-bit of \(c_k\) when \(k\lt i\), the \(X\)-bit of \(c_k\) when \(k=i\), and \(\mathrm{false}\) when \(k\gt i\). Then the family of these \(\{ 0,1\} \times \{ 0,1\} \)-valued pairs, indexed by \(k \in \{ 0,\dots ,n-1\} \), is mutually independent.
Let \(n \ge 1\), and consider vectors of length \(n\) whose entries are drawn from the three-element set \(\{ \text{neither}, \text{leftOnly}, \text{rightOnly}\} \) of disjoint-coordinate labels. Under the uniform law on these \(3^n\) vectors, every individual vector \(c\) carries probability mass
Fix a positive integer \(n\), and consider the finite set of vectors \((x_1,\dots ,x_n)\) of length \(n\) in which each coordinate \(x_i\) takes one of the three values *neither*, *left-only*, or *right-only*. Under the uniform probability law on this set, the total mass — the measure assigned to the set of all such vectors — equals \(1\).
Let \(n \ge 1\) and equip the \(n\)-element set \(\{ 0, 1, \dots , n-1\} \) with the uniform probability law, under which each point carries mass \(1/n\). Then for every subset \(S\) of this set, its probability is \(\left\lvert S\right\rvert /n\), the number of elements of \(S\) divided by \(n\).
Let \(n\) be a positive integer, and let the uniform probability law on the \(n\)-element set \(\{ 0,1,\dots ,n-1\} \) (the finite type \(\mathrm{Fin}\, n\)) assign equal mass to each point. Then for every element \(i\), the mass of the singleton \(\{ i\} \) equals \(\dfrac {1}{n}\).
Let \(n\) be a positive integer, and let \(\mu \) be the uniform probability law on the \(n\)-element set \(\{ 0, 1, \dots , n-1\} \), which assigns to each point equal weight. Then the total mass of \(\mu \) is one; that is, \(\mu (\{ 0,1,\dots ,n-1\} ) = 1\).
Fix \(n \ge 1\) coordinates, and write \(\mathbb {P}\) for the uniform probability measure on the finite space of hard-distribution samples. Let \(E\) denote the disjointness event — the set of samples for which Alice’s input set and Bob’s input set are disjoint — and let \(\mathbb {P}(\, \cdot \mid E)\) be the hard distribution conditioned on disjointness. Then for every event \(A \subseteq E\), conditioning on \(A\) produces the same measure whether one starts from \(\mathbb {P}\) or from \(\mathbb {P}(\, \cdot \mid E)\); that is, \(\mathbb {P}(\, \cdot \mid A) = \mathbb {P}(\, \cdot \mid E)(\, \cdot \mid A)\).
Fix \(n\ge 1\) and let \(\mu \) be the uniform distribution on the hard-sample space, where each sample selects a special coordinate \(T\in \{ 1,\dots ,n\} \), Alice’s and Bob’s bits \(x_T,y_T\) at that coordinate, and a disjoint-coordinate type at every other coordinate; these data determine Alice’s set \(X\) and Bob’s set \(Y\). Write \(\mu [\, \cdot \mid E\, ]\) for \(\mu \) conditioned on an event \(E\), and let \(\nu \) denote \(\mu \) conditioned first on the pair \((X,Y)\) being disjoint and then further on Bob’s special bit satisfying \(y_T=\mathrm{false}\). Then for every set \(S\) of samples,
Fix a positive integer \(n\), and consider the finite space of hard samples, each consisting of a distinguished coordinate \(T \in \{ 0, 1, \dots , n-1\} \), two bits \(x_T, y_T\), and a function assigning to every coordinate one of the three disjointness states (neither, left-only, or right-only); equip this space with its uniform counting measure \(\mathrm{vol}\). The dualization map — which sends \(T\) to its reversed index \(n-1-T\), exchanges the bits \(x_T\) and \(y_T\), and replaces the state function by the one that, at each reversed coordinate, swaps the left-only and right-only states while fixing neither — is measure-preserving for \(\mathrm{vol}\): it is measurable, and the pushforward of \(\mathrm{vol}\) along it is again \(\mathrm{vol}\).
Fix a positive integer \(n\), and consider the finite sample space whose points are the hard samples \(\omega = (T, x_T, y_T, \mathrm{other})\), with \(T \in \{ 0, \dots , n-1\} \), bits \(x_T, y_T \in \{ \text{false}, \text{true}\} \), and a coordinate labelling \(\mathrm{other} \colon \{ 0,\dots ,n-1\} \to \{ \text{neither}, \text{leftOnly}, \text{rightOnly}\} \). Let \(f\) be the self-map of this space that toggles Alice’s special-coordinate bit, sending \(\omega \) to \((T, \lnot x_T, y_T, \mathrm{other})\) while leaving \(T\), \(y_T\), and \(\mathrm{other}\) fixed. Then \(f\) is measure-preserving for the volume (uniform counting) measure on the finite sample space: it is measurable and pushes the volume measure forward to itself.
Fix a positive integer \(n\), and let \(\omega \) be a hard sample for the disjointness lower bound. Writing \(\mathrm{vol}\) for the ambient volume measure on the finite sample space of hard samples, and \(\mathrm{vol}(A)\) for its real-valued mass on a set \(A\), the singleton consisting of the dual sample \(\mathrm{dual}(\omega )\) — obtained by swapping Alice’s and Bob’s roles and reversing the coordinate order — carries the same mass as the singleton \(\{ \omega \} \); that is, \(\mathrm{vol}(\{ \mathrm{dual}(\omega )\} ) = \mathrm{vol}(\{ \omega \} )\).
Fix a positive integer \(n\), and let \(\omega \) be a point of the hard sample space, i.e. a choice of special coordinate \(T \in \{ 0,\dots ,n-1\} \), bits \(X_T, Y_T \in \{ 0,1\} \), and an assignment of one of the three disjointness states to each coordinate. Write \(\omega '\) for the point obtained from \(\omega \) by flipping only Alice’s special-coordinate bit \(X_T\), leaving \(T\), \(Y_T\), and all the other-coordinate data unchanged. Then the ambient volume measure assigns the singleton \(\{ \omega '\} \) the same real-valued mass as the singleton \(\{ \omega \} \).
Work with the volume (counting) measure on the finite space of hard-distribution samples, in which each sample records a special coordinate \(T\), Alice’s and Bob’s special bits \(X_T\) and \(Y_T\), and the remaining per-coordinate data. Fix a bit \(b\) and a triple \(c=(t,u,v)\), where \(t\) is a coordinate and \(u,v\) are bit-vectors, playing the role of a value of the coarse conditioning \((T,\, X_{\lt T},\, Y_{\gt T})\). Then the mass of the event that Bob’s special bit satisfies \(Y_T=\mathrm{false}\), the coarse conditioning equals \(c\), and Alice’s special bit satisfies \(X_T=b\) is exactly \(\tfrac 12\) times the mass of the event that \(Y_T=\mathrm{false}\) and the coarse conditioning equals \(c\).
Fix a positive integer \(n\), and let \(p\) be a deterministic communication protocol whose two inputs are subsets of \(\{ 0,1,\dots ,n-1\} \) — one held by Alice, one by Bob — and whose output is a single bit. Each sample \(\omega \) of the hard distribution determines an input pair, and running \(p\) on that pair records a transcript; together with the sample’s special coordinate \(T\), Alice’s input bits at the coordinates strictly below \(T\), and Bob’s input bits at the coordinates strictly above \(T\), this transcript forms the value \(Z_p(\omega )\). For an achievable value \(z\) of this record, let \(Z_p^{-1}(z)\) be the set of samples \(\omega \) with \(Z_p(\omega )=z\). Let \(p^{\vee }\) be the protocol obtained from \(p\) by exchanging the roles of Alice and Bob and by reversing the coordinate order \(i \mapsto n-1-i\) in both input subsets. Let \(z^{\vee }\) be the value whose special coordinate is \(n-1-T\), whose Alice-bits-before-the-special-coordinate vector is the coordinate reversal of the Bob-bits-after vector recorded by \(z\), whose Bob-bits-after vector is the coordinate reversal of the Alice-bits-before vector recorded by \(z\), and whose transcript is the transcript recorded by \(z\) reinterpreted as a transcript of \(p^{\vee }\) by swapping the two players and re-routing each message through the coordinate reversal \(i \mapsto n-1-i\). Then the volume, taken as a real number, of \(Z_{p^{\vee }}^{-1}(z^{\vee })\) equals the volume, taken as a real number, of \(Z_p^{-1}(z)\).
Fix a positive integer \(n\) and let \(p\) be a deterministic two-party protocol for the length-\(n\) disjointness problem, run on the input pair produced by a sample \(\omega \) drawn from the hard sample space \(\Omega \) (a finite type carrying the volume measure). To each sample associate its \(Z\)-value \(Z_p(\omega )\), which records the transcript of \(p\) on the input of \(\omega \), the special coordinate \(T = T(\omega )\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\); for an achievable value \(z\) write \(Z_p^{-1}(z) \subseteq \Omega \) for the fiber of samples with \(Z_p(\omega ) = z\). Let \(p^{\dagger }\) be the dual protocol, obtained by exchanging the roles of Alice and Bob and reversing the coordinate order, and let \(z^{\dagger }\) be the dual \(Z\)-value of \(z\). Writing \(x_T(\omega )\) and \(y_T(\omega )\) for Alice’s and Bob’s bits at the special coordinate, for every achievable \(Z\)-value \(z\) of \(p\) and every \(b \in \{ 0,1\} \) the real-valued volume measure satisfies
Fix a positive integer \(n\), a deterministic communication protocol \(p\) for the disjointness problem on inputs of length \(n\), and a real number \(\gamma \). For a protocol \(q\) and a hard sample \(\omega \), let \(Z_q(\omega )\) denote the associated statistic (the transcript of \(q\) on the induced input together with the special coordinate and the padded bit vectors lying before and after it), write \(d^X_q(z)\) for the total variation distance between the conditional law of Alice’s special-coordinate bit given \(Z_q = z\) and the uniform law on one bit, and write \(d^Y_q(z)\) for the same distance computed from Bob’s special-coordinate bit. Let \(p^\ast \) be the dual protocol, obtained by exchanging the roles of Alice and Bob and reversing both coordinate sets. Then, on the event that both special-coordinate bits are \(0\), the volume of the set of samples \(\omega \) with \(\gamma \lt d^X_{p^\ast }\big(Z_{p^\ast }(\omega )\big)\) equals the volume of the set of samples \(\omega \) with \(\gamma \lt d^Y_{p}\big(Z_{p}(\omega )\big)\), where the volume is taken in the ambient measure on the finite hard sample space and read as a real number.
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\). On the finite sample space of hard-distribution samples \(\omega \), let \(Z(\omega ) = (M, T, X_{\lt T}, Y_{\gt T})\) be the statistic recording the transcript \(M\) produced by running \(p\) on the input pair generated by \(\omega \), the special coordinate \(T\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\). Then for every achievable value \(z\) of \(Z\)—that is, every \(z\) realized by at least one sample—the fiber \(Z^{-1}(\{ z\} )\), consisting of all samples \(\omega \) with \(Z(\omega ) = z\), has nonzero volume under the ambient measure on the sample space.
Fix a positive integer \(n\), let \(p\) be a deterministic communication protocol for disjointness inputs of length \(n\), and let \(z\) be an achievable value of the \(Z\)-variable of \(p\) — that is, a \(Z\)-value realized by at least one sample of the hard distribution. Form the dual protocol, obtained by exchanging the roles of Alice and Bob and reversing the coordinate order of both input sets, together with the dual \(Z\)-value induced by \(z\). Then, under the ambient measure on the hard-sample space, the fiber of the dual \(Z\)-variable over the dual value of \(z\) (the set of hard samples whose dual \(Z\)-variable equals it) has the same measure as the fiber of the \(Z\)-variable of \(p\) over \(z\).
Let \(n\) be a positive integer and let \(p\) be a deterministic communication protocol for the disjointness inputs of length \(n\). For a sample \(\omega \) of the hard distribution, let \(Z_p(\omega )\) denote the statistic recording the transcript of \(p\) on the input pair generated by \(\omega \), together with the special coordinate \(T\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\) (each padded with \(\mathrm{false}\) off its range); write \(Z_p^{-1}(z)\) for the set of samples on which this statistic takes a fixed achievable value \(z\). Let \(p^{*}\) be the dual protocol, obtained by exchanging the roles of Alice and Bob and reversing both coordinate sets, and let \(z^{*}\) be the correspondingly dualized value of the statistic for \(p^{*}\). Then for each bit \(b\), the set of samples with \(Z_{p^{*}}(\omega ) = z^{*}\) whose special coordinate carries Alice’s bit equal to \(b\) has the same mass, under the ambient measure on the finite space of hard samples, as the set of samples with \(Z_p(\omega ) = z\) whose special coordinate carries Bob’s bit equal to \(b\).
Fix a positive integer \(n\) of coordinates, and let \(\omega \) be any sample of the hard distribution. Then Alice’s input set of the dualized sample coincides with the coordinate-reversal of Bob’s input set of the original sample; that is, \(X(\mathrm{dual}(\omega )) = \mathrm{reverse}(Y(\omega ))\), where reversal sends each coordinate \(i\) to \(n-1-i\).
Let \(\omega _X\) and \(\omega _Y\) be two samples of the hard distribution that agree on the special coordinate \(T\), on Alice’s padded input bits before \(T\), and on Bob’s padded input bits after \(T\). Then the mixed sample, which takes Alice’s data from \(\omega _X\) and Bob’s data from \(\omega _Y\), induces exactly Alice’s input set of \(\omega _X\); that is, \(X(\mathrm{mix}(\omega _X,\omega _Y)) = X(\omega _X)\).
Fix a positive integer \(n\), and let \(\omega \) be a sample of the hard distribution on \(n\) coordinates — a choice of special coordinate \(T\), Alice’s and Bob’s bits there, and a disjointness type at every remaining coordinate. Write \(\omega ^\ast \) for the dual sample obtained by exchanging Alice’s and Bob’s roles and reversing the coordinate order. Then Alice’s before-special vector of \(\omega ^\ast \) equals the coordinate-reversal of Bob’s after-special vector of \(\omega \); concretely, at each coordinate \(i\) Alice’s before-special bit of \(\omega ^\ast \) at \(i\) coincides with Bob’s after-special bit of \(\omega \) at the mirror-image coordinate \(n-1-i\).
Fix an integer \(n \ge 1\) and a deterministic communication protocol \(p\) for disjointness on \(n\) coordinates, and consider the coarse variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) — the transcript \(M\) produced by running \(p\) on the input pair, the special coordinate \(T\), the vector \(X_{\lt T}\) of Alice’s bits strictly below \(T\) (padded with \(\mathrm{false}\) elsewhere), and the vector \(Y_{\gt T}\) of Bob’s bits strictly above \(T\) (padded with \(\mathrm{false}\) elsewhere). If two hard-distribution samples \(\omega \) and \(\omega '\) yield the same value \(z\) of \(Z\), then they determine the same pre-special vector: \(X_{\lt T}(\omega ) = X_{\lt T}(\omega ')\).
Fix a positive integer \(n\) and a deterministic protocol \(p\) for the length-\(n\) disjointness problem, whose inputs are pairs of subsets of \(\{ 1,\dots ,n\} \). Recall that each hard-distribution sample \(\omega \) determines a value \(Z(\omega ) = (M, T, X_{\lt T}, Y_{\gt T})\) of the associated \(Z\) variable, consisting of the transcript, the special coordinate, Alice’s bits strictly before the special coordinate (padded by \(\mathrm{false}\) elsewhere), and Bob’s bits strictly after it. Let \(z\) be any achievable value of this \(Z\) variable. Then for every hard-distribution sample \(\omega \) with \(Z(\omega ) = z\), the function recording Alice’s bits before the special coordinate of \(\omega \) agrees with the \(X_{\lt T}\) component of \(z\).
Let \(n\) be a positive integer and let \(\omega \) be a hard-distribution sample, consisting of a special coordinate \(T \in \{ 0,\dots ,n-1\} \), Alice’s special bit \(x_T\), Bob’s special bit \(y_T\), and a coordinate assignment. Then flipping only Alice’s special bit leaves her padded prefix before the special coordinate unchanged; that is, the function \(i \mapsto \bigl(x\text{-bit of }\omega \text{ at }i\text{ if }i \lt T,\ \text{else }\mathrm{false}\bigr)\) is the same whether it is computed from \(\omega \) or from the sample obtained by toggling \(x_T\) alone.
Let \(\omega _X\) and \(\omega _Y\) be two samples of the hard distribution that share the same special coordinate \(T\), that agree on Alice’s before-special vector (the bits Alice holds at the coordinates strictly below \(T\), padded with \(\mathrm{false}\) elsewhere), and that agree on Bob’s after-special vector (the bits Bob holds at the coordinates strictly above \(T\), padded with \(\mathrm{false}\) elsewhere). Then the mixed sample \(\mathrm{mix}(\omega _X,\omega _Y)\), which takes Alice’s data from \(\omega _X\) and Bob’s data from \(\omega _Y\), has the same before-special vector as \(\omega _X\).
Fix a positive integer \(n\), and let \(\omega \) be a sample of the hard distribution on \(n\) coordinates. For every coordinate \(i \in \{ 0, 1, \dots , n-1\} \), Alice’s bit at coordinate \(i\) in the dualized sample \(\omega ^{\ast }\) coincides with Bob’s bit at the reversed coordinate \(\mathrm{rev}(i) = n-1-i\) in the original sample \(\omega \); that is,
Let \(\omega _X\) and \(\omega _Y\) be two hard-distribution samples that share the same special coordinate \(T\), whose Alice-inputs on the coordinates before \(T\) agree, and whose Bob-inputs on the coordinates after \(T\) agree. Then at every coordinate \(i\), Alice’s bit under the mixed sample \(\mathrm{mix}(\omega _X,\omega _Y)\) equals Alice’s bit under \(\omega _X\) at that coordinate.
Let \(\omega \) be a sample of the hard distribution, with special coordinate \(T\), per-coordinate assignment \(\text{other}\), and let \(\text{xBit}(\omega , i)\) denote Alice’s bit at coordinate \(i\). Then for every coordinate \(i \neq T\), Alice’s bit at \(i\) equals Alice’s bit in the disjoint coordinate-pair \(\text{other}(i)\) that \(\omega \) assigns to \(i\).
Fix a positive integer \(n\), a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), and an achievable value \(z\) of the summary variable \(Z\) associated with \(p\). For any such pair, let the *special-bit distance from uniform* be the total variation distance between the conditional law of the relevant party’s bit at the special coordinate — given that the summary variable equals \(z\) — and the uniform law on a single bit; write it \(d_X\) when the party is Alice and \(d_Y\) when the party is Bob. Passing to the dual protocol \(p^{*}\), obtained by interchanging the roles of Alice and Bob and reversing the coordinate order of both input sets, and to the correspondingly dualized value \(z^{*}\) of \(Z\), one has
That is, Alice’s special-bit distance from uniform in the dual system equals Bob’s special-bit distance from uniform in the original.
Fix a positive integer \(n\), a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), and an achievable value \(z\) of the random variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) for \(p\). Writing \(\mathrm{TV}\) for the total variation distance, the quantity
the total variation distance between the conditional law of Alice’s bit at the special coordinate given \(Z = z\) and the uniform distribution on one bit, is nonnegative.
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for the disjointness function on inputs of length \(n\), and let \(z\) be an achievable value of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), recording the transcript, the special coordinate \(T\), Alice’s bits before \(T\), and Bob’s bits after \(T\). Write \(\mu _z\) for the hard distribution conditioned on the fiber \(\{ Z = z\} \), and let \(U\) be the uniform law on one bit. Suppose the event that Bob’s bit \(Y_T\) at the special coordinate equals \(\mathrm{false}\) has nonzero probability under \(\mu _z\). Then the Kullback–Leibler divergence from \(U\) of the law of Alice’s special bit \(X_T\) under \(\mu _z\) equals the Kullback–Leibler divergence from \(U\) of the law of \(X_T\) under \(\mu _z\) additionally conditioned on the event \(\{ Y_T = \mathrm{false}\} \).
Let \(n\) be a positive integer and let \(p\) be a deterministic communication protocol for the length-\(n\) set-disjointness problem, whose inputs are pairs of subsets of \(\{ 0,\dots ,n-1\} \). Write \(\mathbb {P}\) for the uniform hard distribution on samples \(\omega \) (each specifying a special coordinate together with the coordinatewise data generating Alice’s set \(X(\omega )\) and Bob’s set \(Y(\omega )\)), let \(X_T(\omega )\) and \(Y_T(\omega )\) be Alice’s and Bob’s bits at the special coordinate, and let \(D\) be the event that the generated pair \((X,Y)\) is disjoint. Let \(Z = Z_p\) be the statistic recording the transcript produced by running \(p\) on the generated input, together with the special coordinate, Alice’s input bits strictly before it, and Bob’s input bits strictly after it. Write \(U\) for the uniform law on one bit and \(D_{\mathrm{KL}}\) for Kullback–Leibler divergence. Fix a value \(z\) of \(Z\) and suppose the conditional probability \(\mathbb {P}\big(Y_T = \mathrm{false} \mid Z = z\big)\) is nonzero. Then
both sides taken as real numbers.
Fix \(n \ge 1\) and a deterministic two-party protocol \(p\) for length-\(n\) set-disjointness inputs. To each hard sample \(\omega \) associate the record
consisting of the protocol’s transcript on \(\omega \)’s input, the special coordinate \(T\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\); for a value \(z\) write \(\{ Z = z\} \) for its fiber. Let \(U\) be the uniform law on a single bit, let \(x_T\) denote Alice’s bit at the special coordinate, and let \(D(\mu \, \| \, \nu )\) denote the Kullback–Leibler divergence, read as a nonnegative real number. Let \(\lambda \) be the hard distribution conditioned on the generated input pair being disjoint and on Bob’s special bit \(y_T\) being false. If the fiber \(\{ Z = z\} \) has nonzero mass under \(\lambda \), then the divergence from \(U\) of the law of \(x_T\) under the uniform measure on hard samples conditioned on \(\{ Z = z\} \) equals the divergence from \(U\) of the law of \(x_T\) under \(\lambda \) conditioned on \(\{ Z = z\} \); that is,
where \(\mu _z\) is the law of \(x_T\) under the uniform measure on hard samples restricted to \(\{ Z = z\} \), and \(\lambda _z\) is the law of \(x_T\) under \(\lambda \) restricted to \(\{ Z = z\} \).
Fix a positive integer \(n\), and let \(\omega \) be a sample of the hard distribution on \(n\) coordinates. Form the dual sample by swapping Alice’s and Bob’s roles and reversing the coordinate order. Then Alice’s truncated input vector in the dual sample — her bits up to and including the special coordinate, padded by \(\mathrm{false}\) elsewhere — agrees coordinate-for-coordinate with the coordinate-reversal of Bob’s truncated input vector in \(\omega \), namely his bits from the special coordinate onward, padded by \(\mathrm{false}\) elsewhere.
Fix a positive dimension \(n\) and let \(\omega \) be a hard-distribution sample. Then Alice’s full input vector of the dual sample \(\omega ^{\vee }\) (obtained by swapping the roles of Alice and Bob and reversing the coordinate order) coincides, coordinate by coordinate, with the reversal of Bob’s full input vector of \(\omega \); that is, for every coordinate the value Alice receives in the dual is the value Bob receives in the original at the mirror-image coordinate \(n-1-i\).
Fix a positive integer \(n\), and consider the length-\(n\) set-disjointness problem in which Alice and Bob each hold a subset of \(\{ 0,1,\dots ,n-1\} \). Let \(p\) be a deterministic two-party communication protocol on such inputs with Boolean output, and let \(p^\ast \) be the dual protocol formed from \(p\) by exchanging the roles of Alice and Bob and reversing the coordinate order of both input sets. Draw a random sample \(\omega \) from the hard distribution conditioned on the event that the two generated sets are disjoint; write \(x(\omega ), y(\omega ) \in \{ 0,1\} ^n\) for Alice’s and Bob’s full length-\(n\) input vectors at \(\omega \), and, for a protocol \(q\), write \(\Pi _q(\omega )\) for the transcript produced by running \(q\) on the pair \((x(\omega ), y(\omega ))\). Then, with every information quantity taken over \(\omega \) under the disjointness-conditioned distribution, the conditional mutual information between Alice’s input vector and the transcript of \(p^\ast \) given Bob’s input vector equals the conditional mutual information between Bob’s input vector and the transcript of \(p\) given Alice’s input vector:
Let \(n\) be a positive integer and let \(p\) be a deterministic two‑party communication protocol on disjointness inputs of length \(n\). Sample from the hard distribution conditioned on the event that the two generated sets are disjoint, and write \(X, Y \colon \{ 1,\dots ,n\} \to \{ 0,1\} \) for Alice’s and Bob’s resulting input vectors and \(M\) for the transcript that \(p\) produces on the corresponding input pair. Then the conditional mutual information between \(X\) and \(M\) given \(Y\), computed under this conditioned distribution, satisfies
where \(\mathrm{complexity}(p)\) is the worst‑case number of bits exchanged by \(p\) and \(\log \) denotes the natural logarithm.
Let \(\omega \) be a hard sample on \(n\) coordinates with special coordinate \(T\), and let \(P, Q\) be decidable binary relations on \(\mathrm{Fin}\, n\) such that for every \(i\), \(P(i, \mathrm{rev}(T)) \iff Q(T, \mathrm{rev}(i))\), where \(\mathrm{rev}\) is the order-reversing bijection \(i \mapsto n-1-i\) on \(\mathrm{Fin}\, n\). Then the Boolean vector obtained from Alice’s bits of the dual sample CommunicationComplexity.Functions.Disjointness.RandomizedLowerBound.dualHardSample by keeping \(x_i\) exactly when \(P(i, \mathrm{rev}(T))\) holds (and setting the remaining coordinates to \(\mathtt{false}\)) equals the coordinate reversal of the vector obtained from Bob’s bits of the original sample \(\omega \) by keeping \(y_i\) exactly when \(Q(T, i)\) holds:
Fix \(n \ge 1\), and let \(\omega \) be a sample of the hard distribution on \(n\) coordinates, with Alice’s input set \(X(\omega ) \subseteq \{ 0,\dots ,n-1\} \) and Bob’s input set \(Y(\omega )\). Passing to the dual sample \(\omega '\) — obtained by swapping Alice’s and Bob’s roles and reversing the coordinate order — Bob’s input set of \(\omega '\) is exactly the coordinate reversal of Alice’s input set of \(\omega \):
where \(\operatorname {rev}(i) = n-1-i\). Equivalently, a coordinate \(i\) belongs to Bob’s set for the dualized sample precisely when the reversed coordinate \(\operatorname {rev}(i)\) belongs to Alice’s set for the original sample.
Fix a positive integer \(n\) and consider two samples \(\omega _X\) and \(\omega _Y\) drawn from the hard-distribution sample space on the coordinates \(\{ 0,1,\dots ,n-1\} \). Suppose the two samples share the same special coordinate \(T\), that they agree on Alice’s input bits before the special coordinate, and that they agree on Bob’s input bits after the special coordinate. Then Bob’s input set of the mixed sample \(\operatorname {mix}(\omega _X,\omega _Y)\) — the sample that takes Alice’s data from \(\omega _X\) and Bob’s data from \(\omega _Y\) — coincides with Bob’s input set of \(\omega _Y\).
Fix a coordinate count \(n \ge 1\) and let \(\omega \) be a sample of the hard distribution, with special coordinate \(T\). Form the dual sample \(\omega ^{*}\) by swapping the roles of Alice and Bob and reversing the coordinate order \(i \mapsto \bar\imath \). Then Bob’s after-special vector of \(\omega ^{*}\) — the vector whose entry at each coordinate \(i\) is Bob’s bit of \(\omega ^{*}\) at \(i\) when \(i\) lies strictly after the special coordinate of \(\omega ^{*}\), and is \(\mathrm{false}\) otherwise — equals the coordinate reversal of Alice’s before-special vector of \(\omega \), whose entry at each coordinate \(i\) is Alice’s bit of \(\omega \) at \(i\) when \(i\) lies strictly before \(T\), and is \(\mathrm{false}\) otherwise. Equivalently, for every coordinate \(i\) the \(i\)-th entry of Bob’s after-special vector of \(\omega ^{*}\) agrees with the \(\bar\imath \)-th entry of Alice’s before-special vector of \(\omega \).
Let \(n\) be a positive integer and let \(p\) be a deterministic communication protocol for disjointness inputs of length \(n\). To each sample \(\omega \) of the hard distribution one associates the value \(Z(\omega ) = (M, T, X_{\lt T}, Y_{\gt T})\), recording the transcript \(M\) of \(p\) run on the input generated by \(\omega \), the special coordinate \(T\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\) (each of the last two padded with \(\mathrm{false}\) outside its range). If two samples \(\omega \) and \(\omega '\) yield the same value \(z\) of this variable, then the post-special bit vectors they determine agree: \(Y_{\gt T}(\omega ) = Y_{\gt T}(\omega ')\).
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\). Each hard-distribution sample \(\omega \) determines a value \(Z(\omega ) = (M, T, X_{\lt T}, Y_{\gt T})\), whose four components are the transcript of \(p\) on the input pair induced by \(\omega \), the special coordinate \(T\), Alice’s bits at the coordinates strictly below \(T\) (and false elsewhere), and Bob’s bits at the coordinates strictly above \(T\) (and false elsewhere). If \(z\) is an achievable value of \(Z\) and \(\omega \) is any sample with \(Z(\omega ) = z\), then the bit-string assigning to each coordinate \(i\) Bob’s bit at \(i\) when \(T \lt i\) and false otherwise coincides with the \(Y_{\gt T}\)-component of \(z\).
Let \(n\) be a positive integer, and let \(\omega \) be a sample of the hard distribution, consisting of a distinguished coordinate \(T \in \{ 0,\dots ,n-1\} \), Alice’s special bit \(x_T \in \{ 0,1\} \), Bob’s special bit \(y_T \in \{ 0,1\} \), and an assignment \(i \mapsto c_i\) of a coordinate type to each index \(i\). Let Bob’s bit at coordinate \(i\) be \(y_T\) when \(i = T\) and otherwise the bit determined by the coordinate type \(c_i\), and let \(\beta (\omega ) \in \{ 0,1\} ^n\) be Bob’s post-special vector, whose \(i\)-th entry equals Bob’s bit at \(i\) when \(T \lt i\) and is \(0\) otherwise. Writing \(\omega '\) for the sample obtained from \(\omega \) by replacing \(x_T\) with its negation and leaving \(T\), \(y_T\), and the coordinate-type assignment unchanged, we have
that is, toggling only Alice’s special-coordinate bit leaves Bob’s input bits after the special coordinate unchanged.
Let \(\omega _X\) and \(\omega _Y\) be two samples from the hard distribution on \(n\) coordinates, and suppose they share the same special coordinate, \(T(\omega _X) = T(\omega _Y)\), that Alice’s before-special profiles agree, \(x_{\lt T}(\omega _X) = x_{\lt T}(\omega _Y)\), and that Bob’s after-special profiles agree, \(y_{\gt T}(\omega _X) = y_{\gt T}(\omega _Y)\). Then Bob’s after-special profile of the mixed sample \(\mathrm{mix}(\omega _X,\omega _Y)\) coincides with that of \(\omega _Y\):
Let \(\omega \) be a sample of the hard distribution over \(n\) coordinates, and write \(\mathrm{rev}(i) = n-1-i\) for the reversal of a coordinate \(i \in \{ 0,1,\dots ,n-1\} \). Form the dual sample by swapping Alice’s and Bob’s roles and reversing the coordinate order. Then Bob’s bit at coordinate \(i\) of the dual sample equals Alice’s bit at the reversed coordinate \(\mathrm{rev}(i)\) of the original sample.
Let \(\omega _X\) and \(\omega _Y\) be two samples of the hard distribution that share the same special coordinate \(T\) (so their \(T\)-fields agree), that agree on the profile of Alice’s bits at the coordinates strictly below \(T\) (each padded by \(\mathrm{false}\) at the remaining coordinates), and that agree on the profile of Bob’s bits at the coordinates strictly above \(T\) (again each padded by \(\mathrm{false}\) elsewhere). Form the mixed sample that takes its special coordinate, Alice’s bit there, and Alice’s input from \(\omega _X\) while taking Bob’s bit at \(T\) and Bob’s input from \(\omega _Y\). Then for every coordinate \(i\), Bob’s bit at \(i\) in this mixed sample equals Bob’s bit at \(i\) in \(\omega _Y\).
Fix a positive integer \(n\) and a sample \(\omega \) of the hard distribution, consisting of a distinguished coordinate \(T \in \{ 0, 1, \dots , n-1\} \), two bits \(x_T\) and \(y_T\), and an assignment \(i \mapsto \omega _{\text{other}}(i)\) of a disjoint coordinate-type (one of *neither*, *left only*, *right only*) to each coordinate. For every coordinate \(i \ne T\), Bob’s bit at \(i\) under this sample equals the bit \(y(\omega _{\text{other}}(i))\) read off from the coordinate-type assigned to \(i\); that is, away from the distinguished coordinate the distinguished bit \(y_T\) plays no role.
Let \(n\) be a positive integer and let \(\omega \) be a sample of the hard distribution, with special coordinate \(T\), and write \(\mathrm{dual}(\omega )\) for its dual sample, obtained by exchanging Alice’s and Bob’s bits and reversing the coordinate order. Then Bob’s suffix vector \(y_{\ge T}\) of \(\mathrm{dual}(\omega )\) equals the coordinate-reversal of Alice’s prefix vector \(x_{\le T}\) of \(\omega \): for every coordinate \(i\) with \(0 \le i \le n-1\),
Let \(\omega \) be a sample of the hard distribution, with special coordinate \(T\), and suppose Bob’s bit at the special coordinate is false. Then Bob’s bit-vector from the special coordinate onward — his bits at all coordinates \(i \ge T\), set to false at every coordinate \(i \lt T\) — coincides, as a function \(\mathrm{Fin}\, n \to \{ \text{false}, \text{true}\} \), with Bob’s bit-vector strictly after the special coordinate — his bits at all coordinates \(i \gt T\), set to false elsewhere.
Fix a sample \(\omega \) of the hard distribution, and write \(T\) for its special coordinate. Let \(y_{\ge T}\) be Bob’s truncated input vector, whose value at a coordinate \(i\) is Bob’s bit at \(i\) when \(T \le i\) and is \(\mathrm{false}\) otherwise. Then evaluating this vector at the special coordinate itself gives Bob’s bit at the special coordinate: \(y_{\ge T}(T)\) equals \(y_T\).
Fix a positive integer \(n\) and let \(\omega \) be a hard sample on the \(n\) coordinates. Then Bob’s full input vector of the dualized sample equals the coordinate reversal of Alice’s full input vector of \(\omega \): for every coordinate the value Bob receives after dualizing agrees with the value Alice receives at the mirror-image coordinate of the original sample.
Let \(n\) be a positive integer, and let \(p\) be a deterministic two-party communication protocol for the length-\(n\) set-disjointness problem, in which Alice holds a set \(X \subseteq \{ 0,\dots ,n-1\} \) and Bob holds a set \(Y \subseteq \{ 0,\dots ,n-1\} \). Draw an input pair from the hard distribution conditioned on the disjointness event \(X \cap Y = \varnothing \); write \(x, y \in \{ 0,1\} ^{n}\) for the indicator vectors of \(X\) and \(Y\) respectively, and write \(M\) for the transcript produced by executing \(p\) on the pair \((X, Y)\). The conditional mutual information, measured in nats, between \(y\) and \(M\) given \(x\) then satisfies
where \(\mathrm{complexity}(p)\) denotes the worst-case number of bits exchanged by \(p\).
Fix a positive integer \(n\), a deterministic communication protocol \(p\) for length-\(n\) disjointness inputs, and a value \(z\) of the coarse conditioning variable. Conditioning on the event that this variable equals \(z\) produces a joint law on the special-coordinate bit-pair \((a,b)\) — where \(a\) is Alice’s bit and \(b\) is Bob’s bit at the special coordinate — together with its two one-bit marginals, the conditional law of \(a\) and the conditional law of \(b\). Then the total variation distance between the joint conditional law and the uniform distribution on \(\{ 0,1\} ^2\) is at most the sum of the total variation distances of the two marginals from the uniform distribution on \(\{ 0,1\} \):
Fix a positive integer \(n\) and a deterministic two-party communication protocol \(p\) that outputs a Boolean on pairs of subsets of \([n]\), intended to compute set disjointness. For a sample \(\omega \) of the hard distribution, let \(Z(\omega )\) denote the statistic recording the transcript produced by running \(p\) on the induced input pair \((X_\omega , Y_\omega )\), together with the special coordinate \(T\), Alice’s input bits at coordinates before \(T\), and Bob’s input bits at coordinates after \(T\). Fix one value \(z\) of this statistic, and let \(b\) be the Boolean output attached to the transcript stored in \(z\). Then every sample \(\omega \) with \(Z(\omega ) = z\) whose bits at the special coordinate satisfy \((\text{Alice's special bit}, \text{Bob's special bit}) = (b, b)\) belongs to the protocol error event; that is, the value \(p\) computes on \((X_\omega , Y_\omega )\) differs from the true disjointness value of \(X_\omega \) and \(Y_\omega \).
Fix a positive integer \(n\), a deterministic protocol \(p\) for the disjointness problem on inputs of length \(n\), an achievable value \(z\) of the coarse variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), and a Boolean \(b\). On the sample space of hard samples, write \(\mu \) for the ambient measure, let \(F = \{ Z = z\} \) be the fiber above \(z\), and let \(E\) be the event that Bob’s bit at the special coordinate equals \(b\). Then conditioning the fiber measure \(\mu [\, \cdot \mid F]\) further on \(E\) yields the same measure as conditioning \(\mu \) directly on the intersection \(F \cap E\):
Fix a length \(n \ge 1\) and a deterministic disjointness protocol \(p\) on inputs of length \(n\). Let \(\mu \) be the uniform distribution on the hard sample space, and let \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) be the random variable recording the transcript \(M\) of \(p\) on the generated input pair, the special coordinate \(T\), Alice’s input bits strictly before \(T\), and Bob’s input bits strictly after \(T\); write \(Y_T\) for Bob’s bit at the special coordinate, and call a sample disjoint when Alice’s and Bob’s generated sets are disjoint. Then for every achievable value \(z\) of \(Z\), the measure obtained from \(\mu \) by first conditioning on the fiber \(\{ Z = z\} \) and then further conditioning on \(\{ Y_T = \mathrm{false}\} \) coincides with the measure obtained from \(\mu \) by conditioning on disjointness and on \(\{ Y_T = \mathrm{false}\} \) and then conditioning on \(\{ Z = z\} \).
Fix a length \(n \ge 1\) and a deterministic protocol \(p\) for the disjointness problem on inputs indexed by \(\{ 0,\dots ,n-1\} \), and let \(z\) be an achievable value of the coarse variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) recorded by the protocol. Write \(F_z\) for the fiber of \(Z\) above \(z\), that is, the set of hard-distribution samples \(\omega \) with \(Z(\omega ) = z\), and let \(\mu _z(\cdot ) = \operatorname {vol}(\cdot \mid F_z)\) denote the hard distribution conditioned on the event \(F_z\). Then for every set \(S\) of samples,
where both sides are taken as real numbers. In other words, restricting to the conditioning event \(F_z\) leaves the conditional measure unchanged, so \(\mu _z\) is supported on \(F_z\).
Fix a positive integer \(n\), a deterministic disjointness protocol \(p\) on inputs of length \(n\), and an achievable value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\); let \(F\) denote the fiber of \(Z\) above \(z\), that is, the set of hard-distribution samples \(\omega \) with \(Z(\omega ) = z\). Write \(\mathrm{vol}\) for the ambient volume measure on the (finite) sample space of hard-distribution samples, and let \(\mathrm{vol}(\cdot \mid F)\) be the associated measure conditioned on \(F\). Then for every set \(S\) of hard-distribution samples, the real number assigned to \(S\) by the conditioned measure satisfies
where the factor \(\mathrm{vol}(F)^{-1}\) is the real reciprocal, so that both sides equal \(0\) when \(F\) has measure zero.
Fix a positive integer \(n\), a deterministic protocol \(p\) for disjointness on inputs of length \(n\), and an achievable value \(z\) of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\) recorded during an execution, so that the fiber \(\{ Z = z\} \) is a subset of the sample space of hard instances. Let \(\mu \) denote the hard distribution conditioned both on the generated input pair being disjoint and on Bob’s bit \(Y_T\) at the special coordinate being \(\mathrm{false}\), and let \(\nu \) denote the uniform measure on the sample space conditioned on the fiber \(\{ Z = z\} \). If the \(\mu \)-measure of the fiber \(\{ Z = z\} \) is nonzero, then the \(\nu \)-measure of the event \(\{ Y_T = \mathrm{false}\} \) is nonzero.
Fix a positive integer \(n\) and a deterministic protocol \(p\) for disjointness inputs of length \(n\), and recall that the raw \(Z\)-variable sends each hard-distribution sample \(\omega \) to the tuple \(Z(\omega ) = (M, T, X_{\lt T}, Y_{\gt T})\) consisting of the transcript of \(p\) on \(\omega \)’s input, the special coordinate, and the input bits before and after that coordinate. An achievable \(Z\)-value is a raw value together with a witness that it arises this way. Then for every achievable \(Z\)-value \(z\), there exists a hard-distribution sample \(\omega \) such that \(Z(\omega )\) equals the underlying raw value of \(z\).
Fix a deterministic communication protocol \(p\) for disjointness inputs of length \(n\), and consider achievable values of the variable \(Z = (M, T, X_{\lt T}, Y_{\gt T})\), each carrying a transcript \(M\), a special coordinate \(T \in \{ 0, \dots , n-1\} \), a Boolean string \(X_{\lt T}\) recording Alice’s bits before the special coordinate, and a Boolean string \(Y_{\gt T}\) recording Bob’s bits after it. Then two such achievable values \(z\) and \(z'\) coincide whenever they have the same transcript, the same special coordinate, the same \(X_{\lt T}\), and the same \(Y_{\gt T}\).
Let \(n\) be a positive integer, let \(p\) be a deterministic disjointness protocol on inputs of length \(n\), and let \(\omega \) be a sample from the hard distribution. Forming the \(Z\)-variable — the packaged record consisting of the transcript of \(p\) on the input generated by \(\omega \), together with the special coordinate, Alice’s bits strictly before it, and Bob’s bits strictly after it — commutes with dualization: the \(Z\)-variable of the dual protocol evaluated at the dualized sample coincides with the dualization of the \(Z\)-variable of \(p\) at \(\omega \). That is,
where \(p^\ast \) is the protocol obtained by interchanging Alice and Bob and reversing both coordinate sets, \(\omega ^\ast \) is the sample obtained by swapping Alice’s and Bob’s bits and reversing the coordinate order, and \((\, \cdot \, )^\ast \) on \(Z\)-values is the corresponding recoding.
Fix a positive integer \(n\) and a deterministic communication protocol \(p\) for disjointness inputs of length \(n\). For a sample \(\omega \) of the hard distribution, let \(Z_p(\omega )\) denote its \(Z\)-statistic: the tuple consisting of the transcript of \(p\) on the input pair generated by \(\omega \), the special coordinate of \(\omega \), Alice’s input bits strictly before that coordinate, and Bob’s input bits strictly after it (each bit vector padded with \(\mathrm{false}\) off its range); and for an achievable value \(z\) of this statistic, write its fiber as \(\{ \omega : Z_p(\omega ) = z\} \). Let \(\omega \mapsto \bar\omega \) be the sample duality, which swaps Alice’s and Bob’s roles and reverses the coordinate order, let \(p^\dagger \) be the dual protocol, and let \(z^\dagger \) be the dual value obtained from \(z\) by dualizing its transcript, reversing its special coordinate, and interchanging (with coordinate reversal) its before- and after-special bit vectors. Then a sample \(\omega \) satisfies \(Z_{p^\dagger }(\bar\omega ) = z^\dagger \) if and only if \(Z_p(\omega ) = z\); equivalently, the preimage under sample duality of the fiber of \(Z_{p^\dagger }\) over \(z^\dagger \) is exactly the fiber of \(Z_p\) over \(z\).
Fix a deterministic protocol \(p\) for set disjointness on \(n\) coordinates, and for a sample \(\omega \) of the hard distribution write \(Z(\omega ) = (M, T, X_{\lt T}, Y_{\gt T})\) for the associated variable, where \(M\) is the transcript that \(p\) produces on the input pair generated by \(\omega \), \(T\) is the special coordinate, \(X_{\lt T}\) is Alice’s input restricted to the coordinates before \(T\), and \(Y_{\gt T}\) is Bob’s input restricted to the coordinates after \(T\). Let \(\omega _X\) and \(\omega _Y\) be two samples, and let \(z\) be an achievable value of this variable with \(Z(\omega _X) = z\) and \(Z(\omega _Y) = z\). Then the mixed sample that takes Alice’s data from \(\omega _X\) and Bob’s data from \(\omega _Y\) again satisfies \(Z(\mathrm{mix}(\omega _X, \omega _Y)) = z\).
For a natural number \(n\), let \(\mathrm{equality}_n\) denote the two-party Boolean function on \(n\)-bit strings that returns \(\mathtt{true}\) precisely when Alice’s input \(x \in \mathrm{BoolInput}(n)\) equals Bob’s input \(y \in \mathrm{BoolInput}(n)\). Its deterministic communication complexity, measured in \(\mathbb {N}_\infty \) as the least worst-case number of bits exchanged by any deterministic protocol computing it, is
For every natural number \(n\), the deterministic communication complexity of the \(n\)-bit equality function satisfies
where \(\mathrm{equality}_n : \mathrm{BoolInput}\, n \times \mathrm{BoolInput}\, n \to \mathrm{Bool}\) is the function that returns \(\mathtt{true}\) exactly when Alice’s input equals Bob’s, and the bound is an inequality in \(\mathbb {N}_\infty \).
Let \(\mathrm{equality}_0 \colon \mathrm{BoolInput}\, 0 \times \mathrm{BoolInput}\, 0 \to \mathrm{Bool}\) be the equality function on \(0\)-bit Boolean inputs, and let \(D(\mathrm{equality}_0)\) denote its deterministic communication complexity, that is, the infimum over all deterministic two-party protocols computing it of the worst-case number of bits exchanged, taken in \(\mathbb {N}_\infty \). Then
For a natural number \(n\), CommunicationComplexity.Functions.Equality.equality is the Boolean function \(\mathrm{equality}_n : \mathrm{BoolInput}\, n \times \mathrm{BoolInput}\, n \to \mathrm{Bool}\) that returns true if and only if Alice’s input \(x\) equals Bob’s input \(y\).
\(\mathrm{equalityHashProtocol}(n, k)\) is the standard public-coin protocol for equality in which, given a shared random hash \(h \in \mathrm{HashSpace}(n, k)\), Alice sends \(h(x) \in \mathrm{Fin}(2^k)\) and Bob replies with \(\mathtt{decide}(h(y) = h(x))\).
For all natural numbers \(n\) and \(k\), the public-coin equality hashing protocol on \(n\)-bit Boolean inputs with hash range \(\mathrm{Fin}(2^k)\) — in which, given a shared random hash \(h\), Alice sends the value \(h(x) \in \mathrm{Fin}(2^k)\) and Bob replies with the truth value of whether \(h(y) = h(x)\) — has worst-case communication complexity exactly \(k + 1\) bits.
Fix natural numbers \(n\) and \(k\), let \(x, y \in \mathrm{BoolInput}(n)\) be \(n\)-bit Boolean strings, and let \(h \in \mathrm{HashSpace}(n,k)\) be a hash function sending \(n\)-bit strings into \(\mathrm{Fin}(2^k)\). Running the equality hashing protocol on private inputs \(x\) and \(y\) with \(h\) as the shared public randomness returns the Boolean value \(\mathrm{decide}(h(x) = h(y))\); that is, the protocol outputs \(\mathrm{true}\) precisely when the two hashes \(h(x)\) and \(h(y)\) agree.
For every natural number \(n \ge 1\), the deterministic communication complexity of the \(n\)-bit equality function satisfies
where \(\mathrm{equality}_n\) is the Boolean function on \(\mathrm{BoolInput}\, n \times \mathrm{BoolInput}\, n\) that returns \(\mathtt{true}\) exactly when Alice’s input equals Bob’s input.
For every natural number \(n\) and every real \(\varepsilon \gt 0\), the \(\varepsilon \)-error public-coin randomized communication complexity of the equality function \(\mathrm{equality}_n\) on pairs of \(n\)-bit Boolean strings satisfies
where \(\lceil \varepsilon ^{-1}\rceil \) denotes the least natural number at least \(\varepsilon ^{-1}\) and \(\lceil \, \cdot \, \rceil \) inside the logarithm the ceiling of the base-\(2\) logarithm. That is, there is a public-coin protocol computing \(\mathrm{equality}_n\) with error probability at most \(\varepsilon \) on every input pair whose worst-case number of exchanged bits is bounded by this quantity.
Fix natural numbers \(n\) and \(k\) and a real number \(\varepsilon \) satisfying \(1/2^{k} \lt \varepsilon \). Then the \(\varepsilon \)-error public-coin randomized communication complexity of the equality function \(\mathrm{equality}_n\) on pairs of \(n\)-bit Boolean strings satisfies
where the bound is taken in the extended naturals \(\mathbb {N}\cup \{ \infty \} \).
Let \(\alpha \) be a type, let \(k\) be a natural number, and let \(x, y \in \alpha \). Writing a hash function as a map \(h : \alpha \to \mathrm{Fin}\, k\), the collision event on which \(x\) and \(y\) receive the same hash value decomposes as the union of collision pieces over all possible common values:
Let \(\alpha \) be a finite type and let \(k \ge 1\), so that each hash value lies in \(\{ 0,\dots ,k-1\} \). Consider the space of hash functions \(h \colon \alpha \to \mathrm{Fin}\, k\), equipped with the uniform product probability measure, and fix two distinct points \(x \ne y\) in \(\alpha \). Then the probability that \(x\) and \(y\) collide is at most \(1/k\):
For \(x, y : \alpha \) and \(a : \mathrm{Fin}\, k\), \(\texttt{collisionPiece}\, k\, x\, y\, a\) is the set of hash functions \(h \in \texttt{HashSpace}\, \alpha \, k\) satisfying \(h(x) = a\) and \(h(y) = a\), i.e. the subset of the product space where both \(x\) and \(y\) are mapped to the fixed value \(a\). These pieces partition the collision event \(\{ h \mid h(x) = h(y)\} \).
Let \(\alpha \) be a finite type and let \(k \ge 1\). Fix two distinct elements \(x \ne y\) of \(\alpha \) and a value \(a \in \mathrm{Fin}\, k\), and equip the space of hash functions \(\alpha \to \mathrm{Fin}\, k\) with the product of the uniform probability measures on the coordinates. Then the collision piece consisting of those hash functions \(h\) with \(h(x) = a\) and \(h(y) = a\) has measure
Fix a type \(\alpha \), a natural number \(k\), and two elements \(x, y : \alpha \). For each value \(a \in \{ 0, \dots , k-1\} \), let \(C_a\) denote the set of hash functions \(h : \alpha \to \{ 0, \dots , k-1\} \) with \(h(x) = a\) and \(h(y) = a\). Then the family \((C_a)_a\) is pairwise disjoint: whenever \(a \ne b\), the sets \(C_a\) and \(C_b\) have empty intersection.
Let \(k \ge 1\) be a natural number and let \(a\) be an element of the finite type \(\mathrm{Fin}\, k = \{ 0, 1, \dots , k-1\} \), equipped with its canonical (uniform) volume measure. Then the measure of the singleton \(\{ a\} \) equals \(1/k\), as an element of the extended nonnegative reals \([0, \infty ]\).
Let \(k \ge 1\), and equip the finite set \(\mathrm{Fin}\, k = \{ 0, 1, \dots , k-1\} \) with its uniform probability measure. Then for every element \(a \in \mathrm{Fin}\, k\), the real-valued measure of the singleton \(\{ a\} \) equals \(1/k\).
\(\texttt{HashSpace}\, \alpha \, k\) is the type of all functions \(\alpha \to \mathrm{Fin}\, k\), i.e. the space of hash functions on \(\alpha \) with output range \(\{ 0,\dots ,k-1\} \). Equipped with the product probability measure (uniform over each coordinate), it serves as the ambient probability space for the random-hash argument.
For a deterministic one-way protocol \(p\) for Index on \(n\) bits and a message \(m\) of \(p\), the answer vector \(a(m) \in \{ 0,1\} ^n\) is the string whose \(i\)-th coordinate is Bob’s output \(\mathrm{decode}(m, i)\) on receiving \(m\) with index \(i\).
Alice’s input \(x \in \{ 0,1\} ^n\) is bad for the protocol \(p\) if there is no answer vector \(a \in A\) with \(d_H(x, a) \lt n/4\), where the radius \(n/4\) is the real number (not its floor). Equivalently, every realisable answer vector is at Hamming distance at least \(n/4\) from \(x\).
If \(n \ge 300\) and the deterministic one-way protocol \(p\) for \(\mathrm{IND}_n\) has cost at most \(\lfloor n/10 \rfloor \), then at least \(2^{n-1}\) of the \(2^n\) inputs \(x \in \{ 0,1\} ^n\) are bad for \(p\) (no answer vector of \(p\) lies within Hamming distance \(n/4\) of \(x\)).
The number of incorrect input pairs \((x, i)\) of \(p\) equals the sum, over all strings \(x \in \{ 0,1\} ^n\), of the number of indices \(i\) on which \(p\) errs at \(x\):
The (two-way) deterministic communication complexity of \(\mathrm{IND}_n\) is at most \(\lceil \log _2 n \rceil + 1\):
where \(\lceil \log _2 n \rceil \) is formalised as \(\mathrm{Nat.clog}\, 2\, n\). This is the protocol in which Bob sends his index \(i\) and Alice replies with the single bit \(x_i\).
For a real \(\varepsilon \) and a natural number \(c\), the proposition that every deterministic one-way protocol \(p\) for \(\mathrm{IND}_n\) with \(p.\mathrm{cost} \le c\) has distributional error strictly greater than \(\varepsilon \) under the uniform input distribution CommunicationComplexity.Functions.Indexing.indexingInputDist. This is the hypothesis of the easy direction of Yao’s minimax principle for one-way protocols, specialised to Index.
For \(n \ge 300\), the predicate CommunicationComplexity.Functions.Indexing.deterministicOneWayDistributionalLowerBound holds at error \(\varepsilon = 1/9\) and cost budget \(c = \lfloor n/10 \rfloor \): every deterministic one-way protocol for \(\mathrm{IND}_n\) of cost at most \(\lfloor n/10 \rfloor \) has uniform-distributional error greater than \(1/9\).
Let \(p\) be a deterministic one-way protocol for \(\mathrm{IND}_n\). If at least \(2^{n-1}\) of Alice’s \(2^n\) inputs are bad for \(p\) (no answer vector of \(p\) lies within Hamming distance \(n/4\)), then the distributional error of \(p\) on \(\mathrm{IND}_n\) under the uniform distribution on pairs \((x, i)\) is at least \(1/8\).
If every deterministic one-way protocol for \(\mathrm{IND}_n\) of cost at most \(\lfloor n/10 \rfloor \) has uniform-distributional error greater than \(1/8\), then the public-coin one-way communication complexity of \(\mathrm{IND}_n\) at error \(1/8\) is strictly greater than \(\lfloor n/10 \rfloor \). This is the previous theorem specialised to the source’s constants \(\varepsilon = 1/8\), \(c = \lfloor n/10 \rfloor \).
For \(n \ge 300\), the public-coin one-way communication complexity of \(\mathrm{IND}_n\) at error \(1/9\) is strictly greater than \(\lfloor n/10 \rfloor \):
This is the Kremer–Nisan–Ron linear lower bound with explicit constants.
Let \(\mathrm{good}\) be the set of inputs \(x \in \{ 0,1\} ^n\) that are not bad for \(p\) (those within Hamming distance \(\lt n/4\) of some answer vector). Then
where \(A\) is the answer set of \(p\).
For every \(n\) and every centre \(a \in \{ 0,1\} ^n\), the Hamming ball of radius \(\lfloor n/4 \rfloor \) around \(a\) satisfies
(as real numbers).
The hard input distribution for Index: the uniform probability measure on pairs \((x, i)\) in \(\mathrm{BoolInput}\, n \times \mathrm{Fin}\, n\), so that Alice’s string \(x\) and Bob’s index \(i\) are chosen independently and uniformly at random. It is packaged as a finite probability space whose measure is \(\mathrm{uniformOn}\, \mathrm{Set.univ}\).
If every deterministic one-way protocol for \(\mathrm{IND}_n\) of cost at most \(c\) has uniform-distributional error greater than \(\varepsilon \), then the public-coin one-way communication complexity of \(\mathrm{IND}_n\) at error \(\varepsilon \) is strictly greater than \(c\) (as an extended natural number). This is the easy direction of Yao’s minimax principle for one-way protocols, applied to Index.
For a fixed input string \(x\), the number of indices \(i\) on which \(p\) errs equals the Hamming distance between \(x\) and the answer vector of Alice’s message on \(x\):
For \(n \ge 300\), every deterministic one-way protocol \(p\) for \(\mathrm{IND}_n\) with \(p.\mathrm{cost} \le \lfloor n/10 \rfloor \) has distributional error strictly greater than \(1/9\) under the uniform distribution on pairs \((x, i)\).
The one-way protocol for Index in which Alice sends her entire string \(x\) as the message (the message type is \(\mathrm{Fin}\, n \to \mathrm{Bool}\) and the send map is the identity), and Bob outputs the \(i\)-th bit of the received message.
Fix a natural number \(n\) and equip the set \(\{ 0,1\} ^n \times \{ 0,1\} ^n\) of pairs of \(n\)-bit Boolean vectors with the uniform probability measure. Let \(\mathrm{IP}_n(x,y) = \bigoplus _{i=1}^{n}(x_i \wedge y_i)\) be the mod-2 inner product. Then for every rectangle \(R \subseteq \{ 0,1\} ^n \times \{ 0,1\} ^n\) — that is, every set of the form \(R = A \times B\) with \(A, B \subseteq \{ 0,1\} ^n\) — the discrepancy of \(\mathrm{IP}_n\) on \(R\) satisfies
where the discrepancy is the expectation over a uniformly random pair \((x,y)\) of \(\mathbf{1}_R(x,y)\, \sigma (\mathrm{IP}_n(x,y))\), with \(\sigma \) the sign map sending \(\mathrm{false} \mapsto 1\) and \(\mathrm{true} \mapsto -1\).
Fix \(n \in \mathbb {N}\), and for \(u,v \in \mathrm{BoolInput}\, n\) write \(\langle u, v\rangle \in \mathrm{Bool}\) for the mod-\(2\) inner product, and \(s(b) \in \{ \pm 1\} \) for the Boolean sign of \(b\) (so \(s(\mathrm{false}) = 1\) and \(s(\mathrm{true}) = -1\)). Then for all \(x, y, z \in \mathrm{BoolInput}\, n\),
where \(y \oplus z\) denotes the coordinatewise XOR of \(y\) and \(z\).
Fix \(n\), work over the type \(\mathrm{BoolInput}(n)\) of \(n\)-bit Boolean inputs equipped with a finite probability measure on the product \(\mathrm{BoolInput}(n)\times \mathrm{BoolInput}(n)\), and let \(\mathrm{IP}\) denote the mod-\(2\) inner product. For any two subsets \(A, B \subseteq \mathrm{BoolInput}(n)\), the discrepancy of \(\mathrm{IP}\) on the combinatorial rectangle \(A\times B\) factors as the iterated integral
where \(\mathbf{1}_A\) and \(\mathbf{1}_B\) are the indicator functions of \(A\) and \(B\), and \(\sigma \) is the Boolean sign map sending \(\mathrm{false}\mapsto 1\) and \(\mathrm{true}\mapsto -1\).
Let \(B \subseteq \mathrm{BoolInput}(n)\) be a set of \(n\)-bit Boolean inputs, and let \(\mathbf{1}_B \colon \mathrm{BoolInput}(n) \to \mathbb {R}\) denote its indicator function, taking the value \(1\) on \(B\) and \(0\) elsewhere. Then for every \(y \in \mathrm{BoolInput}(n)\) one has \(\mathbf{1}_B(y) \le 1\).
Let \(B\) be a set of \(n\)-bit Boolean inputs, and let \(\mathbf{1}_B \colon (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathbb {R}\) be its indicator function, taking the value \(1\) at points of \(B\) and \(0\) elsewhere. Then for every \(n\)-bit Boolean input \(y\),
\(\mathrm{innerProduct}(n, x, y) : \mathrm{Bool}\) is the mod-2 inner product of two \(n\)-bit vectors \(x\) and \(y\): it returns true exactly when the number of coordinates where both inputs are true is odd, i.e. when \(\bigoplus _{i=1}^{n} (x_i \wedge y_i) = 1\).
Let \(x, y\) be \(n\)-bit Boolean inputs, and let \(i\) be a coordinate at which \(y_i = \mathtt{true}\). Write \(\langle u, v\rangle = \bigoplus _{j} (u_j \wedge v_j)\) for the mod-2 inner product of two \(n\)-bit inputs, and let \(x'\) denote the input agreeing with \(x\) at every coordinate except \(i\), where \(x'_i = \lnot \, x_i\). Then
so negating the \(i\)-th coordinate of \(x\) flips the value of the inner product.
Let \(f \colon \mathrm{BoolInput}\, n \to \mathbb {R}\) be any real-valued function on the \(n\)-bit Boolean inputs. Then, integrating \(f\) against the uniform probability measure on the \(2^n\) inputs,
so that the integral equals the arithmetic average of the values \(f(x)\) taken over all \(x \in \mathrm{BoolInput}\, n\).
Let \(B \subseteq \{ 0,1\} ^n\) be an arbitrary set of \(n\)-bit Boolean inputs, with indicator \(\mathbf1_B\), and write \(\langle x,y\rangle = \bigoplus _{i=1}^{n} (x_i \wedge y_i)\) for the mod-2 inner product of \(x,y \in \{ 0,1\} ^n\). Taking \(x\) and \(y\) to range independently over \(\{ 0,1\} ^n\) under the uniform probability measure,
Here \((-1)^{\langle x,y\rangle }\) is the \(\pm 1\) value that is \(-1\) when the inner product is \(1\) and \(+1\) when it is \(0\); equivalently, the average over uniformly random \(x\) of the square of the average over uniformly random \(y\) of \(\mathbf1_B(y)\, (-1)^{\langle x,y\rangle }\) is at most \(2^{-n}\).
Let \(n\) be a natural number, and regard the type of \(n\)-bit Boolean inputs \(\mathrm{Fin}\, n \to \mathrm{Bool}\) as a finite probability space, with its associated probability mass function \(p\). Then for every input \(x\), the mass assigned to \(x\), viewed as a real number, is
Let \(k, n \in \mathbb {N}\) and let \(\varepsilon \in \mathbb {R}\). If
then \(k\) is strictly smaller than the \(\varepsilon \)-error public-coin randomized communication complexity of the mod-2 inner product function on \(n\)-bit inputs; that is, \(k \lt \mathrm{R}^{\mathrm{pub}}_{\varepsilon }\! \bigl(\mathrm{IP}_{n}\bigr)\).
Fix \(n \in \mathbb {N}\) and equip the product \(\mathrm{BoolInput}(n) \times \mathrm{BoolInput}(n)\) with its uniform probability measure. For any two sets \(A, B \subseteq \mathrm{BoolInput}(n)\), the discrepancy of the mod-2 inner product \(\mathrm{IP}_n\) over the combinatorial rectangle \(A \times B\) satisfies
Fix \(n \in \mathbb {N}\), and let \(z\) be an \(n\)-bit Boolean input that is \(\mathrm{true}\) in at least one coordinate, say \(z_i = \mathrm{true}\) for some index \(i\). For a Boolean input \(x\), write \(\langle x, z\rangle = \bigoplus _{j} (x_j \wedge z_j)\) for their mod-2 inner product, and let \(s(x) \in \{ \pm 1\} \) be the sign \(+1\) when \(\langle x, z\rangle \) is \(\mathrm{false}\) and \(-1\) when it is \(\mathrm{true}\). Then summing this sign over all \(2^n\) Boolean inputs \(x\) gives
Fix a bit-string \(z \in \{ 0,1\} ^n\), and for each \(x \in \{ 0,1\} ^n\) let \(\langle x, z\rangle \in \{ 0,1\} \) denote the mod-2 inner product of \(x\) and \(z\). Summing the Boolean sign \((-1)^{\langle x, z\rangle }\) over all \(2^n\) inputs \(x\) gives
Fix \(n \in \mathbb {N}\). For a Boolean value \(b\), let \(\varepsilon (b) \in \{ \pm 1\} \) denote its Boolean sign, equal to \(+1\) when \(b\) is \(\mathrm{false}\) and \(-1\) when \(b\) is \(\mathrm{true}\); and for \(n\)-bit Boolean inputs \(x, y\), let \(\langle x, y\rangle \) denote their mod-2 inner product (which is \(\mathrm{true}\) exactly when the number of coordinates \(i\) with \(x_i = y_i = \mathrm{true}\) is odd). Then for any two \(n\)-bit Boolean inputs \(y\) and \(z\),
where the sum ranges over all \(2^n\) Boolean inputs \(x\).
Let \(\alpha \) be a finite type with decidable equality, let \(h \colon \alpha \times \alpha \to \mathbb {R}\) be any function, and let \(c \in \mathbb {R}\). Summing \(h(y,z)\) against the weight that equals \(c\) when \(y = z\) and \(0\) otherwise, over all ordered pairs, leaves only the diagonal terms:
Let \(\alpha \) and \(\beta \) be finite types and let \(f : \alpha \to \beta \to \mathbb {R}\). Then
That is, squaring each row-sum of \(f\) and summing over rows yields the same value as summing, over every ordered pair of columns \((y,z)\), the inner product of those two columns.
Let \(n \in \mathbb {N}\), and identify the set of \(n\)-bit Boolean inputs with \(\{ 0,1\} ^n\). For \(x, y \in \{ 0,1\} ^n\), write \(\langle x, y\rangle = \bigoplus _{i=1}^{n} (x_i \wedge y_i) \in \{ 0,1\} \) for their mod-\(2\) inner product, so that \((-1)^{\langle x, y\rangle }\) equals \(+1\) when \(\langle x, y\rangle = 0\) and \(-1\) when \(\langle x, y\rangle = 1\). Then for every subset \(B \subseteq \{ 0,1\} ^n\),
Let \(b\) be an arbitrary real-valued function on the set of \(n\)-bit Boolean inputs. For two such inputs \(x,y\), write \(\langle x,y\rangle = \sum _{i=1}^{n}(x_i \wedge y_i)\) for the number of coordinates on which both are true, so that \((-1)^{\langle x,y\rangle }\) is the sign of their mod-2 inner product, equal to \(+1\) when \(\langle x,y\rangle \) is even and \(-1\) when it is odd. Then
where both the inner and outer sums range over all \(n\)-bit Boolean inputs.
Let \(\alpha \) and \(\beta \) be finite index sets, and let \(\phi : \alpha \times \beta \to \mathbb {R}\) be a real-valued array whose columns are mutually orthogonal with a common squared length: for a fixed constant \(c \in \mathbb {R}\),
for all \(y, z \in \beta \). Then for any weights \(b : \beta \to \mathbb {R}\),
Let \(y, z \colon \mathrm{Fin}\, n \to \mathrm{Bool}\) be two \(n\)-bit Boolean inputs, and let \(i \in \mathrm{Fin}\, n\) be a coordinate. Then the \(i\)-th coordinate of their coordinatewise xor equals the xor of their individual \(i\)-th coordinates, namely \(y_i \oplus z_i\).
Let \(y, z \in \mathbb {F}_2^n\) be two \(n\)-bit Boolean strings, and let \(y \oplus z\) denote their coordinatewise xor, the string whose \(i\)-th bit is \(y_i \oplus z_i\). Then \(y \oplus z\) is the all-zero string if and only if \(y = z\).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \) with \(\mu \) absolutely continuous with respect to \(\nu \), and write \(f = \frac{d\mu }{d\nu }\) for the (real-valued) Radon–Nikodym density. Suppose the log-likelihood ratio \(x \mapsto \log \frac{d\mu }{d\nu }(x)\) is \(\mu \)-integrable. Then
where the left-hand inner integral is the \(L^1(\nu )\)-distance of the density from \(1\) and the right-hand integrand is the Kullback–Leibler integrand \(t \mapsto t\log t - t + 1\).
Let \((\Omega , \mu )\) be a probability space, and let \(\varphi (t) = t \log t - t + 1\) for \(t \ge 0\) denote the convex generator of the Kullback–Leibler divergence. Let \(f : \Omega \to \mathbb {R}\) be a measurable, integrable, nonnegative function with \(\int _\Omega f \, d\mu = 1\), so that \(f\) is the density of a probability measure with respect to \(\mu \), and assume that \(x \mapsto \varphi (f(x))\) is integrable with respect to \(\mu \). Then
Let \(\alpha \) be a finite alphabet, and consider words of length \(n\) over \(\alpha \). For any centre word \(u \in \alpha ^n\) and any radius \(r\), the Hamming ball \(B_r(u)\) has cardinality equal to the ball volume at radius \(r\) for a length-\(n\) code over an alphabet of size \(\left\lvert \alpha \right\rvert \); that is,
Let \(u \in \alpha ^n\) be a word of length \(n\) over a finite alphabet \(\alpha \), and let \(r\) be a natural number. Then the Hamming ball of radius \(r\) centred at \(u\) is the disjoint union of the Hamming spheres of radii \(0, 1, \dots , r\) centred at \(u\):
Let \(\alpha \) be a finite alphabet and fix a word \(u \in \alpha ^n\) of length \(n\), where the Hamming distance \(d(x,y)\) between two words counts the coordinates in which they differ. For every radius \(k \in \mathbb {N}\), the Hamming sphere \(S_k(u) = \{ v \in \alpha ^n : d(u,v) = k\} \) of words at Hamming distance exactly \(k\) from \(u\) has cardinality
Let \(\alpha \) be a finite alphabet, let \(u \in \alpha ^n\) be a word of length \(n\), and let \(r\) be a natural number. Then the Hamming sphere \(S_r(u)\) of radius \(r\) centred at \(u\) is the union, taken over all \(r\)-element subsets \(S\) of the \(n\) coordinate positions, of the support fibers of \(u\) over \(S\) — that is, the sets of words differing from \(u\) at exactly the positions in \(S\):
Let \(\alpha \) be a finite alphabet, fix a centre word \(u \in \alpha ^n\), and let \(r\) and \(t\) be natural numbers with \(r \ne t\). Then the Hamming sphere \(S_r(u) = \{ v \in \alpha ^n : d(u,v) = r\} \) of radius \(r\) and the Hamming sphere \(S_t(u) = \{ v \in \alpha ^n : d(u,v) = t\} \) of radius \(t\), both centred at \(u\), are disjoint, where \(d\) denotes the Hamming distance.
Let \(\alpha \) be a finite alphabet, let \(u \in \alpha ^n\) be a word of length \(n\), and let \(r\) be a natural number. For \(t \in \mathbb {N}\), write \(S_t(u)\) for the Hamming sphere of radius \(t\) centred at \(u\). Then the family \(\{ S_t(u)\} _{t \in \{ 0,1,\ldots ,r-1\} }\) is pairwise disjoint: for any two distinct indices \(s, t\) with \(s, t \lt r\), the spheres \(S_s(u)\) and \(S_t(u)\) share no common word.
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\frac{d\mu }{d\nu } : \Omega \to \mathbb {R}\) be the real-valued Radon–Nikodym density of \(\mu \) with respect to \(\nu \) (the extended nonnegative Radon–Nikodym derivative converted to a real number). Then the function \(x \mapsto \left\lvert \frac{d\mu }{d\nu }(x) - 1\right\rvert \) is integrable with respect to \(\nu \).
Let \((\Omega ,\mu )\) be a measure space, let \(X\colon \Omega \to \mathbb {R}\) be a measurable function, and let \(t\in \mathbb {R}\). Write \(\Lambda (t)\) for the cumulant generating function of \(X\) under \(\mu \), that is, the logarithm of the moment generating function evaluated at \(t\). If the function \(x\mapsto e^{t\, X(x)}\) is \(\mu \)-integrable, then so is the function \(x\mapsto e^{t\, X(x)-\Lambda (t)}\).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(f\colon \Omega \to \mathbb {R}\) be the real-valued Radon–Nikodym density of \(\mu \) with respect to \(\nu \), that is, the function whose value at each point \(x\) is the real number obtained from the Radon–Nikodym derivative \(\frac{d\mu }{d\nu }(x)\). Then \(f\) is integrable with respect to \(\nu \).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\frac{d\mu }{d\nu }\colon \Omega \to \mathbb {R}\) denote the real-valued Radon–Nikodym density of \(\mu \) with respect to \(\nu \). Then the function \(x \mapsto \frac{d\mu }{d\nu }(x) - 1\) is integrable with respect to \(\nu \).
Let \(\mu \) be a probability measure on a measurable space \(\Omega \), let \(X : \Omega \to \mathbb {R}\) be a random variable, and let \(t \in \mathbb {R}\). Write \(\Lambda (t) = \log \int _\Omega e^{t X(x)}\, d\mu (x)\) for the cumulant generating function of \(X\) at \(t\). If the function \(x \mapsto e^{t X(x)}\) is \(\mu \)-integrable, then
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \) with \(\mu \) absolutely continuous with respect to \(\nu \), and write \(\frac{d\mu }{d\nu }\) for the (real-valued) Radon–Nikodym density of \(\mu \) with respect to \(\nu \). Assume that the log-likelihood ratio \(x \mapsto \log \frac{d\mu }{d\nu }(x)\) is \(\mu \)-integrable. Then
where \(\varphi (t) = t\log t + 1 - t\).
Let \(\mu \) be a finite measure on a measurable space \(\Omega \), and let \(g : \Omega \to \mathbb {R}\) be measurable and integrable with mean zero, \(\int _\Omega g\, d\mu = 0\). Then the integral of \(g\) over the region where it is nonnegative equals half its \(L^1\) norm:
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and suppose that \(\mu \) is absolutely continuous with respect to \(\nu \), written \(\mu \ll \nu \). Then the Radon–Nikodym derivative \(\frac{d\mu }{d\nu }\) integrates to one against \(\nu \):
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and suppose that \(\mu \) is absolutely continuous with respect to \(\nu \). Writing \(\frac{d\mu }{d\nu }\) for the Radon–Nikodym density of \(\mu \) with respect to \(\nu \) (the real-valued version of the Radon–Nikodym derivative), one has
Given a probability mass function \(p\) on \(\mathrm{Fin}\, m\) and a natural number \(n\), \(\mathrm{cdf}(p, n)\) is the cumulative probability \(\sum _{j \lt n} p(j)\), computed as an extended non-negative real (\(\mathbb {R}_{\ge 0}^\infty \)).
Let \(p\) be a probability mass function on \(\mathrm{Fin}\, m\), and for each natural number \(n\) let \(\mathrm{cdf}(p, n) = \sum _{j \lt n} p(j) \in \mathbb {R}_{\ge 0}^\infty \) be the total mass assigned to the indices below \(n\). Then the map \(n \mapsto \mathrm{cdf}(p, n)\) is monotone: whenever \(n_1 \le n_2\), one has \(\mathrm{cdf}(p, n_1) \le \mathrm{cdf}(p, n_2)\).
Let \(p\) be a probability mass function on the finite set \(\mathrm{Fin}\, m = \{ 0, 1, \dots , m-1\} \). Then its cumulative distribution function evaluated at \(m\) accounts for the entire mass, \(\mathrm{cdf}(p, m) = 1\), this value being computed in the extended non-negative reals \(\mathbb {R}_{\ge 0}^\infty \).
Let \(p\) be a probability mass function on \(\mathrm{Fin}\, m\), and let \(n \in \mathrm{Fin}\, m\). Then the cumulative distribution function satisfies \(\mathrm{cdf}(p, n+1) = \mathrm{cdf}(p, n) + p(n)\), where the sums are taken as extended non-negative reals in \(\mathbb {R}_{\ge 0}^\infty \).
Let \((f_i)_{i \in \iota }\) be a family of values in the extended natural numbers \(\mathbb {N}_\infty = \mathbb {N}\cup \{ \infty \} \), indexed by an arbitrary type \(\iota \), and let \(n \in \mathbb {N}\). Then \(\inf _{i} f_i \le n\) if and only if there exists an index \(i\) with \(f_i \le n\), where \(n\) is regarded as an element of \(\mathbb {N}_\infty \) via the canonical inclusion.
Given a PMF \(p\) on \(\mathrm{Fin}\, m\) (with \(m \ge 1\)) and a value \(x \in \mathbb {R}_{\ge 0}^\infty \), \(\mathrm{invCdf}(p, x)\) is the largest index \(i \in \mathrm{Fin}\, m\) satisfying \(\mathrm{cdf}(p, i) \le x\), i.e. the generalized inverse (quantile) of the CDF.
Let \(p\) be a probability mass function on \(\mathrm{Fin}\, m = \{ 0, 1, \dots , m-1\} \) with \(m \ge 1\), and write \(F(n) = \sum _{j \lt n} p(j) \in \mathbb {R}_{\ge 0}^{\infty }\) for its cumulative distribution function. Let \(x \in \mathbb {R}_{\ge 0}^{\infty }\) satisfy \(x \lt 1\), and let \(i \in \mathrm{Fin}\, m\). Then the largest index whose cumulative probability does not exceed \(x\) equals \(i\) if and only if
Let \(\Omega _X\) and \(\Omega _Y\) be finite probability spaces, and let \(\delta \gt 0\). Give each coin-tape space \(\mathrm{CoinTape}(n)\) the uniform probability measure arising from \(n\) independent fair coin flips (mass \(2^{-n}\) on each string), and give every product space — both \(\Omega _X \times \Omega _Y\) and \(\mathrm{CoinTape}(n_X) \times \mathrm{CoinTape}(n_Y)\) — the corresponding product measure. Then there exist \(n_X, n_Y \in \mathbb {N}\) and maps \(\varphi _X : \mathrm{CoinTape}(n_X) \to \Omega _X\) and \(\varphi _Y : \mathrm{CoinTape}(n_Y) \to \Omega _Y\) such that, writing \(\varphi _X \times \varphi _Y\) for the map \((u,v) \mapsto (\varphi _X(u), \varphi _Y(v))\), every subset \(S \subseteq \Omega _X \times \Omega _Y\) satisfies
Let \(\Omega \) be a finite probability space, and let \(\delta \gt 0\). Then there exist a natural number \(n\) and a map \(\varphi : \mathrm{CoinTape}(n) \to \Omega \) from the length-\(n\) binary strings such that, writing \(\mu _n\) for the uniform probability measure on \(\mathrm{CoinTape}(n)\) (each of the \(2^n\) strings equally likely) and \(\mu \) for the probability measure on \(\Omega \), every subset \(S \subseteq \Omega \) satisfies
Given a PMF \(p\) on \(\mathrm{Fin}\, m\) and a positive integer \(n\), the map \(\mathrm{uniformApprox}(p, n) : \mathrm{Fin}\, n \to \mathrm{Fin}\, m\) sends \(j \mapsto \mathrm{invCdf}(p, j/n)\), discretizing the unit interval into \(n\) equally spaced points and mapping each through the quantile function of \(p\).
Let \(p\) be a probability mass function on \(\mathrm{Fin}\, m\) with \(m \ge 1\), and let \(n \ge 1\) be an integer. For each \(j \in \mathrm{Fin}\, n\), let \(\varphi (j) \in \mathrm{Fin}\, m\) be the largest index \(k\) whose cumulative probability \(\sum _{\ell \lt k} p(\ell )\) does not exceed \(j/n\); equivalently, \(\varphi \) evaluates the generalized inverse (quantile function) of the cumulative distribution function of \(p\) at the \(n\) equally spaced points \(0, 1/n, \dots , (n-1)/n\). Then for every index \(i \in \mathrm{Fin}\, m\), the fraction of these \(n\) points that \(\varphi \) sends to \(i\) exceeds the true probability \(p(i)\) by at most \(1/n\):
where \(p(i)\) is read as a real number.
Let \(\alpha \) be a finite type, let \(p, q, g : \alpha \to \mathbb {R}\), and suppose \(0 \le g(a) \le 1\) for every \(a \in \alpha \). Fix a real number \(\delta \), and suppose that for every subset \(T \subseteq \alpha \) one has \(\sum _{a \in T} p(a) \le \sum _{a \in T} q(a) + \delta \). Then \(\sum _{a} p(a)\, g(a) \le \sum _{a} q(a)\, g(a) + \delta \).
Let \(\Omega \) be a finite measurable space, and let \(\mu \) and \(\nu \) be finite measures on \(\Omega \) with \(\mu \) absolutely continuous with respect to \(\nu \). Write \(\tfrac {d\mu }{d\nu }\) for the Radon–Nikodym derivative and let \(\mathrm{llr}(\omega ) = \log \tfrac {d\mu }{d\nu }(\omega )\) be the log-likelihood ratio at \(\omega \). Then the Kullback–Leibler divergence satisfies
where \([\, x\, ]_{+}\) denotes the value in \([0,\infty ]\) obtained by viewing the real number \(x\) as a nonnegative extended real, negative values being sent to \(0\).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and consider the real-valued Radon–Nikodym density \(x \mapsto \left(\tfrac {d\mu }{d\nu }(x)\right)\), obtained from the \([0,\infty ]\)-valued Radon–Nikodym derivative \(\tfrac {d\mu }{d\nu }\colon \Omega \to [0,\infty ]\) by taking its real value at each point. Then this function \(\Omega \to \mathbb {R}\) is measurable.
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and for each \(x \in \Omega \) let \(\frac{d\mu }{d\nu }(x) \in \mathbb {R}\) denote the Radon–Nikodym derivative of \(\mu \) with respect to \(\nu \), taken as a real number. Then the set \(\left\{ x \in \Omega : \frac{d\mu }{d\nu }(x) \ge 1\right\} \), on which \(\mu \) locally dominates \(\nu \), is a measurable subset of \(\Omega \).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \) with \(\mu \) absolutely continuous with respect to \(\nu \) (that is, \(\mu \ll \nu \)), and let \(S \subseteq \Omega \). Writing \(\frac{d\mu }{d\nu }\) for the (real-valued) Radon–Nikodym derivative of \(\mu \) with respect to \(\nu \), one has
Let \(u\) be a word of length \(n\) over a finite alphabet \(\alpha \), let \(S\) be a set of coordinate positions, and let \(i\) be a position with \(i \notin S\). For every symbol \(a \in \alpha \) with \(a \ne u_i\), the words in the support fiber of \(u\) over \(\{ i\} \cup S\) — those differing from \(u\) at exactly the positions in \(\{ i\} \cup S\) — whose entry at position \(i\) equals \(a\) are exactly as many as the words in the support fiber of \(u\) over \(S\), that is, the words differing from \(u\) at exactly the positions in \(S\).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\mathrm{TV}(\mu ,\nu )\) denote their total variation distance and \(\mathrm{KL}(\mu \, \| \, \nu )\) their Kullback–Leibler divergence. Then
where the left-hand side is regarded as an element of the extended nonnegative reals \([0,\infty ]\).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and suppose that \(\mu \) is absolutely continuous with respect to \(\nu \). Then the total variation distance \(\mathrm{TV}(\mu ,\nu )\) and the Kullback–Leibler divergence \(D_{\mathrm{KL}}(\mu \, \| \, \nu )\) (both taking values in \([0,\infty ]\)) satisfy
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and suppose \(\mu \) is absolutely continuous with respect to \(\nu \) and that the log-likelihood ratio \(\log \frac{d\mu }{d\nu }\) is \(\mu \)-integrable. Then, as an inequality in the extended nonnegative reals \(\mathbb {R}_{\ge 0} \cup \{ \infty \} \),
where \(\mathrm{TV}(\mu ,\nu )\) is the total variation distance between \(\mu \) and \(\nu \) and \(\mathrm{KL}(\mu \, \| \, \nu )\) is the Kullback–Leibler divergence of \(\mu \) from \(\nu \).
The \(\varepsilon \)-error private-coin randomized communication complexity of \(f : X \to Y \to \alpha \) is the extended natural number
i.e. the minimum worst-case number of bits exchanged over all private-coin randomized protocols that compute \(f\) with error at most \(\varepsilon \) on every input.
Let \(f : X \times Y \to \alpha \) be a function of two arguments, and let \(\varepsilon \ge 0\) be an error parameter. Then the \(\varepsilon \)-error private-coin randomized communication complexity of \(f\) is at most its deterministic communication complexity:
as an inequality of extended natural numbers.
Let \(f : X \to Y \to \alpha \) be a function, let \(\varepsilon \in \mathbb {R}\) be an error bound, and let \(n \in \mathbb {N}\). Then the \(\varepsilon \)-error private-coin randomized communication complexity of \(f\) is at most \(n\) if and only if there exist coin-tape lengths \(n_X, n_Y \in \mathbb {N}\) and a private-coin protocol \(p\) with Alice’s randomness space \(\{ 0,1\} ^{n_X}\) and Bob’s randomness space \(\{ 0,1\} ^{n_Y}\), inputs in \(X\) and \(Y\), and outputs in \(\alpha \), such that \(p\) computes \(f\) with error at most \(\varepsilon \) on every input and the communication complexity of \(p\) is at most \(n\).
Let \(f : X \to Y \to \alpha \) be a function, let \(\varepsilon \in \mathbb {R}\) be an error bound, and let \(n \in \mathbb {N}\). Then the \(\varepsilon \)-error private-coin randomized communication complexity satisfies \(R_\varepsilon (f) \le n\) if and only if there exist coin-tape lengths \(n_X, n_Y \in \mathbb {N}\) and a private-coin finite-message protocol \(p\) with randomness spaces \(\{ 0,1\} ^{n_X}\) and \(\{ 0,1\} ^{n_Y}\) (over input types \(X\), \(Y\) and output type \(\alpha \)) such that \(p\) computes \(f\) with error at most \(\varepsilon \) on every input pair and has worst-case communication cost at most \(n\). Thus the private-coin complexity, defined via binary-message protocols, is unchanged when one instead allows protocols whose players may send elements of arbitrary finite nonempty message alphabets, each \(\beta \)-valued message being charged \(\lceil \log _2 \left\lvert \beta \right\rvert \rceil \) bits.
Let \(X\), \(Y\), and \(\alpha \) be types, let \(f : X \to Y \to \alpha \) be the function to be computed, and let \(\Omega _X\) and \(\Omega _Y\) be finite probability spaces serving as the two players’ private randomness. Fix real numbers \(\varepsilon '\) and \(\varepsilon \) with \(\varepsilon ' \lt \varepsilon \), and let \(p\) be a finite-message private-coin protocol with randomness spaces \(\Omega _X\) and \(\Omega _Y\), input types \(X\) and \(Y\), and output type \(\alpha \), and suppose that \(p\) computes \(f\) with error at most \(\varepsilon '\), i.e. for every input pair \((x,y)\) the probability that \(p\) returns a value other than \(f(x,y)\) is at most \(\varepsilon '\). Then the \(\varepsilon \)-error private-coin randomized communication complexity of \(f\) satisfies \(R_\varepsilon (f) \le \mathrm{complexity}(p)\), where the right-hand side is the worst-case number of bits exchanged by \(p\).
Let \(f : X \to Y \to \alpha \) be a function of two arguments, and let \(R_\varepsilon (f)\) denote its \(\varepsilon \)-error private-coin randomized communication complexity — the least worst-case number of bits exchanged by any private-coin protocol computing \(f\) with probability of error at most \(\varepsilon \) on every input, taken as an extended natural number. Then for any two error bounds \(\varepsilon ', \varepsilon \in \mathbb {R}\) with \(\varepsilon ' \le \varepsilon \), we have \(R_\varepsilon (f) \le R_{\varepsilon '}(f)\); that is, allowing a larger error can only decrease or preserve the required communication.
A private-coin finite-message protocol over input types \(X\), \(Y\) with private randomness types \(\Omega _X\), \(\Omega _Y\) and output type \(\alpha \) is defined as a deterministic finite-message protocol whose effective input for Alice is \(\Omega _X \times X\) and for Bob is \(\Omega _Y \times Y\).
Given a function \(f : X \to \Omega _X \to \beta \) (Alice’s message function depending on her public input and private randomness) and a continuation \(P : \beta \to \mathsf{Protocol}\), constructs a protocol node at which Alice sends a \(\beta \)-valued message and execution branches on it.
A protocol \(p\) \(\varepsilon \)-computes a function \(f : X \to Y \to \alpha \) if for every input pair \((x, y)\),
Let \(p\) be a private-coin finite-message protocol over input types \(X\), \(Y\) with private-randomness types \(\Omega _X\), \(\Omega _Y\) and output type \(\alpha \), each of \(\Omega _X\) and \(\Omega _Y\) carrying a measure, let \(f : X \times Y \to \alpha \) be a target function, and let \(\varepsilon \in \mathbb {R}\). Then the proposition that \(p\) \(\varepsilon \)-computes \(f\) coincides with the proposition that \(p\) \(\varepsilon \)-satisfies the predicate \(Q(x,y,a)\) given by \(a = f(x,y)\); that is, the two propositions are equal.
A protocol \(p\) \(\varepsilon \)-satisfies a predicate \(Q : X \to Y \to \alpha \to \mathrm{Prop}\) if for every input pair \((x, y)\),
Given a protocol \(p\) with output type \(\alpha \) and a family \(q : \alpha \to \mathrm{Protocol}\; \Omega _X\; \Omega _Y\; X\; Y\; \beta \), CommunicationComplexity.Deterministic.FiniteMessage.Protocol.bind \(p\) \(q\) is the protocol that runs \(p\), observes its output \(a\), and then executes \(q\, a\) on the same inputs using the same randomness spaces.
Let \(p\) be a private-coin finite-message protocol with input types \(X\), \(Y\), private-coin types \(\Omega _X\), \(\Omega _Y\), and output type \(\alpha \), and let \(q\) assign to each value \(a \in \alpha \) a private-coin finite-message protocol \(q(a)\) over input types \(X\), \(Y\) and private-coin types \(\Omega _X\), \(\Omega _Y\) with output type \(\beta \). Then for all inputs \(x \in X\), \(y \in Y\) and private coins \(\omega _x \in \Omega _X\), \(\omega _y \in \Omega _Y\),
where \(a \in \alpha \) is the output of executing \(p\) on inputs \(x, y\) with coins \(\omega _x, \omega _y\). That is, running the sequential composition of \(p\) and \(q\) on inputs \(x, y\) with private coins \(\omega _x, \omega _y\) gives the value obtained by first running \(p\) on \(x, y\) with coins \(\omega _x, \omega _y\) to produce an output \(a\), and then running \(q(a)\) on \(x, y\) with coins \(\omega _x, \omega _y\).
Given a function \(f : Y \to \Omega _Y \to \beta \) (Bob’s message function depending on his public input and private randomness) and a continuation \(P : \beta \to \mathsf{Protocol}\), constructs a protocol node at which Bob sends a \(\beta \)-valued message and execution branches on it.
Given functions \(hX : \Omega _X' \to \Omega _X\) and \(hY : \Omega _Y' \to \Omega _Y\), CommunicationComplexity.PrivateCoin.FiniteMessage.Protocol.comapRandomness \(hX\) \(hY\) \(p\) reindexes the randomness spaces of \(p\), yielding a protocol over \(\Omega _X'\) and \(\Omega _Y'\) that maps coins through \(hX\) and \(hY\) before passing them to \(p\).
Let \(p\) be a private-coin finite-message protocol over input types \(X\), \(Y\) with private-randomness types \(\Omega _X\), \(\Omega _Y\) and output type \(\alpha \), and let \(hX : \Omega _X' \to \Omega _X\) and \(hY : \Omega _Y' \to \Omega _Y\) be arbitrary maps of randomness spaces. Reindexing the randomness of \(p\) through \(hX\) and \(hY\) — that is, forming the protocol over \(\Omega _X'\), \(\Omega _Y'\) that maps its coins through \(hX\) and \(hY\) before feeding them to \(p\) — yields a protocol with the same worst-case communication complexity as \(p\).
Let \(p\) be a private-coin finite-message protocol over input types \(X\), \(Y\) with private-randomness types \(\Omega _X\), \(\Omega _Y\) and output type \(\alpha \), and let \(hX : \Omega _X' \to \Omega _X\) and \(hY : \Omega _Y' \to \Omega _Y\) be functions reindexing the two randomness spaces. Then for all inputs \(x : X\), \(y : Y\) and all coins \(\omega _x : \Omega _X'\), \(\omega _y : \Omega _Y'\), executing the reindexed protocol on \(x, y\) with coins \(\omega _x, \omega _y\) yields the same output as executing \(p\) itself on \(x, y\) with the pushed-forward coins \(hX(\omega _x)\) and \(hY(\omega _y)\).
Given a function \(g : \alpha \to \beta \) and a private-coin protocol \(p\) over randomness spaces \(\Omega _X, \Omega _Y\) with output type \(\alpha \), map \(g\) \(p\) is the protocol that runs \(p\) and then applies \(g\) to its output, yielding a protocol with output type \(\beta \) and the same communication tree.
Let \(p\) be a private-coin finite-message protocol with input types \(X\), \(Y\), private-randomness types \(\Omega _X\), \(\Omega _Y\), and output type \(\alpha \), and let \(g \colon \alpha \to \beta \) be any function. Then the protocol obtained by running \(p\) and applying \(g\) to its output has the same communication complexity as \(p\), i.e. \(\mathrm{complexity}(\mathrm{map}\, g\, p) = \mathrm{complexity}(p)\).
Let \(p\) be a private-coin finite-message communication protocol with input types \(X\) and \(Y\), private-randomness types \(\Omega _X\) and \(\Omega _Y\), and output type \(\alpha \), and let \(g \colon \alpha \to \beta \). Write \(\mathrm{map}(g, p)\) for the protocol that runs \(p\) and then applies \(g\) to the result, and let executing a protocol on inputs \(x, y\) with private coins \(\omega _x, \omega _y\) denote running it on the paired inputs \((\omega _x, x)\) and \((\omega _y, y)\). Then for all \(x \in X\), \(y \in Y\), \(\omega _x \in \Omega _X\), and \(\omega _y \in \Omega _Y\),
Let \(p\) be a binary private-coin protocol with randomness spaces \(\Omega _X\) and \(\Omega _Y\), input spaces \(X\) and \(Y\), and output type \(\alpha \). Then the private-coin finite-message protocol obtained by embedding \(p\) has the same communication complexity as \(p\); that is, viewing each Boolean message of \(p\) as an element of the two-element message alphabet leaves the worst-case number of bits exchanged unchanged.
Let \(p\) be a binary private-coin protocol with randomness spaces \(\Omega _X\) and \(\Omega _Y\), input spaces \(X\) and \(Y\), and output type \(\alpha \), in which Alice’s and Bob’s messages are single bits. Then there exists a private-coin finite-message protocol \(P\) over the same spaces such that, for every pair of inputs \(x \in X\), \(y \in Y\) and every pair of private coins \(\omega _x \in \Omega _X\), \(\omega _y \in \Omega _Y\), the two protocols return the same output when run on \(x, y\) with coins \(\omega _x, \omega _y\), and moreover \(P\) has exactly the same communication complexity as \(p\).
The embedding of a binary private-coin protocol into the finite-message framework does not change what the protocol computes. Let \(p\) be a binary private-coin communication protocol with input spaces \(X\) and \(Y\), private-randomness spaces \(\Omega _X\) and \(\Omega _Y\), and output type \(\alpha \), and let \(\widehat{p}\) be the finite-message private-coin protocol obtained by embedding \(p\) (interpreting each Boolean message as an element of a two-element message alphabet). Then for every pair of inputs \(x \in X\), \(y \in Y\) and every pair of private coins \(\omega _x \in \Omega _X\), \(\omega _y \in \Omega _Y\), the randomized execution of \(\widehat{p}\) agrees with that of \(p\):
Given a family of protocols \(p : (i : \mathrm{Fin}\, k) \to \mathrm{Protocol}\; (\Omega _{Xf}\, i)\; (\Omega _{Yf}\, i)\; X\; Y\; (\alpha _f\, i)\) with heterogeneous randomness and output types, pi \(p\) runs all \(k\) protocols in parallel on the same inputs using independent randomness from the product spaces \((\prod _i \Omega _{Xf}\, i,\, \prod _i \Omega _{Yf}\, i)\), returning the tuple of outputs \((i : \mathrm{Fin}\, k) \to \alpha _f\, i\).
Let \(k\) be a natural number and let \(p_0, \dots , p_{k-1}\) be a family of private-coin finite-message protocols indexed by \(i \in \mathrm{Fin}\, k\), where \(p_i\) has private-randomness types \(\Omega _{X,i}\), \(\Omega _{Y,i}\), shared input types \(X\), \(Y\), and output type \(\alpha _i\). Then the communication complexity of their \(k\)-fold product protocol equals the sum of the individual communication complexities,
Let \(k\) be a natural number and, for each \(i \in \mathrm{Fin}\, k\), let \(p_i\) be a private-coin finite-message protocol with private-randomness types \(\Omega _{X,i}\) for Alice and \(\Omega _{Y,i}\) for Bob, common public input types \(X\) and \(Y\), and output type \(\alpha _i\). Fix public inputs \(x \in X\) and \(y \in Y\), together with coin vectors \(\omega _x \in \prod _{i} \Omega _{X,i}\) and \(\omega _y \in \prod _{i} \Omega _{Y,i}\). Then executing the product protocol on these data agrees componentwise with executing each factor on its own coordinate of the randomness:
Given protocols \(p_1\) and \(p_2\) with potentially distinct randomness spaces \((\Omega _{X_1}, \Omega _{Y_1})\) and \((\Omega _{X_2}, \Omega _{Y_2})\), CommunicationComplexity.Deterministic.FiniteMessage.Protocol.prod \(p_1\) \(p_2\) runs them in parallel on the same inputs using independent randomness drawn from \(\Omega _{X_1} \times \Omega _{X_2}\) and \(\Omega _{Y_1} \times \Omega _{Y_2}\), returning the pair of their outputs in \(\alpha _1 \times \alpha _2\).
Let \(p_1\) and \(p_2\) be private-coin finite-message protocols, both over the same input types \(X\) and \(Y\), with private-randomness types \((\Omega _{X_1}, \Omega _{Y_1})\) and \((\Omega _{X_2}, \Omega _{Y_2})\) and output types \(\alpha _1\) and \(\alpha _2\) respectively. Their product, the protocol on inputs \(X\) and \(Y\) that runs \(p_1\) and \(p_2\) in parallel using independent randomness drawn from \(\Omega _{X_1} \times \Omega _{X_2}\) and \(\Omega _{Y_1} \times \Omega _{Y_2}\) and returns the pair of their outputs, has worst-case communication cost (in bits) equal to the sum of the individual costs:
Let \(p_1\) and \(p_2\) be private-coin finite-message protocols over the same input types \(X\) and \(Y\), with output types \(\alpha _1\) and \(\alpha _2\) and private-randomness types \((\Omega _{X_1}, \Omega _{Y_1})\) and \((\Omega _{X_2}, \Omega _{Y_2})\) respectively. Then for all inputs \(x : X\) and \(y : Y\), and all paired private coins \(\omega _x = (\omega _{x,1}, \omega _{x,2}) \in \Omega _{X_1} \times \Omega _{X_2}\) and \(\omega _y = (\omega _{y,1}, \omega _{y,2}) \in \Omega _{Y_1} \times \Omega _{Y_2}\), executing the product protocol of \(p_1\) and \(p_2\) yields the pair
where each component is the output of the corresponding factor run on the same inputs with its own share of the private randomness.
Given a private-coin protocol \(p_1\) and a deterministic protocol \(p_2\), CommunicationComplexity.PrivateCoin.FiniteMessage.Protocol.prodDet \(p_1\) \(p_2\) runs both on the same inputs and pairs their outputs in \(\alpha _1 \times \alpha _2\), with \(p_2\) contributing no additional randomness.
Fix input types \(X\), \(Y\), a private-randomness type \(\Omega _X\) for Alice and \(\Omega _Y\) for Bob, and output types \(\alpha _1\), \(\alpha _2\). Let \(p_1\) be a private-coin finite-message protocol producing an output in \(\alpha _1\), and let \(p_2\) be a deterministic finite-message protocol over \(X\), \(Y\) producing an output in \(\alpha _2\). Form their product, the private-coin protocol that runs \(p_1\) and then \(p_2\) on the same inputs (with \(p_2\) using no private randomness) and reports the pair of their outputs in \(\alpha _1 \times \alpha _2\). Then the worst-case communication cost of this product equals the sum of the individual costs, that is, the complexity of the product is the complexity of \(p_1\) plus the complexity of \(p_2\).
Let \(X\), \(Y\) be input types and \(\Omega _X\), \(\Omega _Y\) private-randomness types. Let \(p_1\) be a private-coin finite-message protocol over these types with output type \(\alpha _1\), and let \(p_2\) be a deterministic finite-message protocol over \(X\), \(Y\) with output type \(\alpha _2\). Then for all inputs \(x : X\), \(y : Y\) and all private coins \(\omega _x : \Omega _X\), \(\omega _y : \Omega _Y\), executing the product protocol of \(p_1\) and \(p_2\) on \(x\), \(y\) with coins \(\omega _x\), \(\omega _y\) yields the pair whose first component is the output of \(p_1\) executed on \(x\), \(y\) with coins \(\omega _x\), \(\omega _y\) and whose second component is the output of the deterministic protocol \(p_2\) run on \(x\), \(y\).
Given a protocol \(p\) over randomness \((\Omega _X, \Omega _Y)\) and a family \(q : \alpha \to \mathrm{Protocol}\; \Omega _X'\; \Omega _Y'\; X\; Y\; \beta \), CommunicationComplexity.PrivateCoin.FiniteMessage.Protocol.rbind \(p\) \(q\) runs \(p\) with the first component of the coin pair, obtains its output \(a\), then runs \(q\, a\) using a fresh, independent second component, yielding a protocol over \((\Omega _X \times \Omega _X', \Omega _Y \times \Omega _Y')\).
Let \(p\) be a private-coin finite-message protocol over inputs \(X\), \(Y\) with private randomness \((\Omega _X, \Omega _Y)\) and output type \(\alpha \), and let \(q\) assign to each output \(a \in \alpha \) a private-coin finite-message protocol \(q\, a\) over \(X\), \(Y\) with fresh randomness \((\Omega _X', \Omega _Y')\) and output type \(\beta \). Suppose there is a natural number \(c\) with \(\mathrm{complexity}(q\, a) = c\) for every \(a \in \alpha \). Then the protocol obtained by running \(p\) and then \(q\, a\) with independent randomness satisfies
Let \(p\) be a private-coin finite-message protocol over inputs \(X\), \(Y\) with private randomness types \(\Omega _X\), \(\Omega _Y\) and output type \(\alpha \), and let \(q\) assign to each output \(a\in \alpha \) a private-coin finite-message protocol \(q(a)\) over the same inputs \(X\), \(Y\) with fresh randomness types \(\Omega _X'\), \(\Omega _Y'\) and output type \(\beta \). Then for all inputs \(x\in X\), \(y\in Y\) and all coin pairs \(\omega _x=(\omega _x^1,\omega _x^2)\in \Omega _X\times \Omega _X'\) and \(\omega _y=(\omega _y^1,\omega _y^2)\in \Omega _Y\times \Omega _Y'\), executing the fresh-randomness bind of \(p\) and \(q\) agrees with running \(p\) on the first coordinates of the coins to obtain an output \(a\) and then running \(q(a)\) on the second coordinates:
\(\texttt{rrun}\, p\, x\, y\, \omega _x\, \omega _y\) executes the protocol \(p\) on public inputs \(x : X\), \(y : Y\) with Alice’s private coin \(\omega _x : \Omega _X\) and Bob’s private coin \(\omega _y : \Omega _Y\), returning an element of \(\alpha \). It is defined by running the underlying deterministic protocol on the paired inputs \((\omega _x, x)\) and \((\omega _y, y)\).
Let \(p\) be a private-coin finite-message protocol over input types \(X\), \(Y\), with private-randomness types \(\Omega _X\), \(\Omega _Y\) and output type \(\alpha \). For all inputs \(x : X\) and \(y : Y\) and all private coins \(\omega _x : \Omega _X\) and \(\omega _y : \Omega _Y\), executing \(p\) with those explicit coins agrees with executing the underlying deterministic finite-message protocol on the paired inputs \((\omega _x, x)\) for Alice and \((\omega _y, y)\) for Bob; that is, both yield the same element of \(\alpha \).
Given a finite-message private-coin protocol \(p\) over finite probability spaces \(\Omega _X\) and \(\Omega _Y\), and a tolerance \(\delta \gt 0\), \(\mathrm{toCoinTape}(p, \delta )\) produces \(n_X\), \(n_Y\), and a new finite-message protocol over \(\mathrm{CoinTape}(n_X) \times \mathrm{CoinTape}(n_Y)\) obtained by precomposing \(p\) with the inverse-CDF approximating maps \(\varphi _X\) and \(\varphi _Y\).
Let \(p\) be a private-coin finite-message protocol over input types \(X\), \(Y\) and output type \(\alpha \), whose private randomness for Alice and Bob is drawn from finite probability spaces \(\Omega _X\) and \(\Omega _Y\). Let \(Q : X \to Y \to \alpha \to \mathrm{Prop}\) be a predicate on input–output triples, and let \(\varepsilon \) and \(\delta \) be real numbers with \(\delta \gt 0\). Suppose \(p\) \(\varepsilon \)-satisfies \(Q\); that is, for every input pair \((x, y)\) the probability, over Alice’s and Bob’s private coins \((\omega _x, \omega _y)\), that the output of \(p\) fails to satisfy \(Q\, x\, y\) is at most \(\varepsilon \). Then the protocol obtained by converting \(p\) at tolerance \(\delta \) — replacing the private randomness by fair coin tapes of some lengths \(n_X\) and \(n_Y\) — \((\varepsilon + \delta )\)-satisfies \(Q\).
Let \(\Omega _X\) and \(\Omega _Y\) be finite probability spaces, let \(X\), \(Y\) be input types and \(\alpha \) an output type, and let \(p\) be a private-coin finite-message protocol over \(\Omega _X\), \(\Omega _Y\) (that is, a deterministic finite-message protocol on the input types \(\Omega _X \times X\) and \(\Omega _Y \times Y\)). For every real tolerance \(\delta \gt 0\), the coin-tape protocol produced from \(p\) and \(\delta \) — the finite-message protocol over \(\mathrm{CoinTape}(n_X) \times X\) and \(\mathrm{CoinTape}(n_Y) \times Y\) obtained by precomposing \(p\) with the inverse-CDF approximating maps — has communication complexity exactly equal to that of \(p\).
Let \(p\) be a private-coin finite-message protocol with private randomness spaces \(\Omega _X\) and \(\Omega _Y\), input types \(X\) and \(Y\), and output type \(\alpha \), in which each message is an element of an arbitrary finite nonempty type \(\beta \) costing \(\lceil \log _2 \left\lvert \beta \right\rvert \rceil \) bits. Converting \(p\) into the associated binary private-coin protocol, in which every message is encoded as a string of bits, leaves the worst-case communication complexity unchanged: the complexity of the converted binary protocol equals the complexity of \(p\).
Let \(p\) be a private-coin finite-message protocol with private-randomness spaces \(\Omega _X\) and \(\Omega _Y\), input spaces \(X\) and \(Y\), and output type \(\alpha \), in which each message is an element of an arbitrary finite nonempty type. Fix inputs \(x \in X\) and \(y \in Y\) together with private coins \(\omega _x \in \Omega _X\) and \(\omega _y \in \Omega _Y\). Then converting \(p\) to a binary private-coin protocol leaves its randomized execution unchanged: running the converted protocol on \(x, y\) with coins \(\omega _x, \omega _y\) returns the same element of \(\alpha \) as running \(p\) itself on the same inputs and coins.
Let \(f : X \to Y \to \alpha \) be a function, let \(\varepsilon \in \mathbb {R}\), and let \(R_\varepsilon (f) \in \mathbb {N}\cup \{ \infty \} \) denote its \(\varepsilon \)-error private-coin randomized communication complexity, namely the least worst-case number of bits exchanged by any private-coin protocol that computes \(f\) with error probability at most \(\varepsilon \) on every input. Then for every \(n \in \mathbb {N}\), the inequality \(n \le R_\varepsilon (f)\) holds (in the extended natural numbers) if and only if every private-coin protocol \(p\) — over coin tapes \(\{ 0,1\} ^{n_X}\) for Alice and \(\{ 0,1\} ^{n_Y}\) for Bob, for arbitrary lengths \(n_X, n_Y \in \mathbb {N}\) — that \(\varepsilon \)-computes \(f\) has communication complexity at least \(n\).
A private-coin protocol with randomness spaces \(\Omega _X\) and \(\Omega _Y\), input spaces \(X\) and \(Y\), and output type \(\alpha \) is defined as a deterministic protocol \(\texttt{Deterministic.Protocol}\; (\Omega _X \times X)\; (\Omega _Y \times Y)\; \alpha \). Alice’s message function receives the pair \((\omega _x, x)\) and Bob’s receives \((\omega _y, y)\), so each player’s coin flip is invisible to the other.
Given a function \(f : X \to \Omega _X \to \mathrm{Bool}\) and a continuation \(P : \mathrm{Bool} \to \mathrm{Protocol}\; \Omega _X\; \Omega _Y\; X\; Y\; \alpha \), constructs the protocol node at which Alice sends the bit \(f\, x\, \omega _x\) and execution continues with \(P(\text{bit})\).
A private-coin protocol \(p\) \(\varepsilon \)-computes a function \(f : X \to Y \to \alpha \) if for every input pair \((x, y)\),
This formalises the standard notion of bounded-error private-coin randomized communication complexity.
Let \(p\) be a private-coin protocol with randomness spaces \(\Omega _X\) and \(\Omega _Y\), input spaces \(X\) and \(Y\), and output type \(\alpha \), each randomness space carrying a measure, let \(f : X \to Y \to \alpha \) be a target function, and let \(\varepsilon \in \mathbb {R}\). Then the proposition that \(p\) approximately computes \(f\) within error \(\varepsilon \) coincides with the proposition that \(p\) approximately satisfies, within error \(\varepsilon \), the predicate that holds of \((x, y, a)\) exactly when \(a = f(x, y)\).
A private-coin protocol \(p\) \(\varepsilon \)-satisfies a predicate \(Q : X \to Y \to \alpha \to \mathrm{Prop}\) if for every input pair \((x, y)\),
where \(\mu \) is the product measure on \(\Omega _X \times \Omega _Y\).
Given a function \(f : Y \to \Omega _Y \to \mathrm{Bool}\) and a continuation \(P : \mathrm{Bool} \to \mathrm{Protocol}\; \Omega _X\; \Omega _Y\; X\; Y\; \alpha \), constructs the protocol node at which Bob sends the bit \(f\, y\, \omega _y\) and execution continues with \(P(\text{bit})\).
\(\texttt{rrun}\; p\; x\; y\; \omega _x\; \omega _y\) executes the private-coin protocol \(p\) on inputs \(x \in X\) and \(y \in Y\) with Alice’s private coin \(\omega _x \in \Omega _X\) and Bob’s private coin \(\omega _y \in \Omega _Y\), returning the output \(\alpha \). It is defined by \(p.\mathrm{run}\, (\omega _x, x)\, (\omega _y, y)\).
Let \(p\) be a private-coin protocol with randomness spaces \(\Omega _X\) and \(\Omega _Y\), input spaces \(X\) and \(Y\), and output type \(\alpha \). Then for all inputs \(x \in X\) and \(y \in Y\) and all coin outcomes \(\omega _x \in \Omega _X\) and \(\omega _y \in \Omega _Y\), the randomized execution of \(p\) equals the deterministic execution of \(p\) on the paired inputs \((\omega _x, x)\) and \((\omega _y, y)\):
The \(\varepsilon \)-error public-coin randomized communication complexity of \(f : X \to Y \to \alpha \) is defined as
the infimum (in \(\mathbb {N}_\infty \)) of the worst-case bit complexity over all public-coin randomized protocols that compute \(f\) with error at most \(\varepsilon \) on every input.
Let \(f : X \to Y \to \alpha \) be a function of Alice’s input and Bob’s input, let \(\varepsilon \in \mathbb {R}\) be an error parameter, and let \(m \in \mathbb {N}\). Then the \(\varepsilon \)-error public-coin communication complexity satisfies \(\mathrm{R}^{\mathrm{pub}}_\varepsilon (f) \le m\) if and only if there exist a length \(n \in \mathbb {N}\) and a public-coin protocol \(p\) over the coin tape \(\mathrm{CoinTape}(n)\) such that \(p\) \(\varepsilon \)-computes \(f\) and the worst-case bit complexity of \(p\) is at most \(m\).
Fix an output type \(\alpha \), a function \(f : X \to Y \to \alpha \), an error tolerance \(\varepsilon \in \mathbb {R}\), and a natural number \(m\). Then the \(\varepsilon \)-error public-coin randomized communication complexity of \(f\) satisfies \(\mathrm{R}^{\mathrm{pub}}_\varepsilon (f) \le m\) if and only if there exist a coin-tape length \(n \in \mathbb {N}\) and a public-coin finite-message protocol \(p\) over the shared randomness \(\mathrm{CoinTape}(n)\) such that \(p\) \(\varepsilon \)-computes \(f\) — that is, for every input pair \((x,y)\) the probability over the coin tape \(\omega \) that \(p\) outputs a value other than \(f(x,y)\) is at most \(\varepsilon \) — and the communication complexity of \(p\) is at most \(m\).
Let \(X\), \(Y\), and \(\alpha \) be types, let \(\Omega \) be a finite probability space, and let \(f : X \to Y \to \alpha \). Fix real numbers \(\varepsilon ' \lt \varepsilon \), and suppose \(p\) is a public-coin finite-message protocol with shared randomness \(\Omega \) and private inputs in \(X\) and \(Y\) that \(\varepsilon '\)-computes \(f\), meaning that for every input pair \((x,y)\) the probability over \(\omega \in \Omega \) that the protocol’s output differs from \(f(x,y)\) is at most \(\varepsilon '\). Then the \(\varepsilon \)-error public-coin randomized communication complexity of \(f\) satisfies
where \(\mathrm{complexity}(p)\) is the worst-case bit cost of \(p\) (a \(\beta \)-valued message costing \(\lceil \log _2 \left\lvert \beta \right\rvert \rceil \) bits) and the inequality holds in \(\mathbb {N}\cup \{ \infty \} \).
Let \(f : X \to Y \to \alpha \) be a function of two arguments, and let \(\varepsilon , \varepsilon ' \in \mathbb {R}\) be error parameters with \(\varepsilon ' \le \varepsilon \). Then the \(\varepsilon \)-error public-coin randomized communication complexity of \(f\) is at most its \(\varepsilon '\)-error public-coin randomized communication complexity: \(\mathrm{R}^{\mathrm{pub}}_\varepsilon (f) \le \mathrm{R}^{\mathrm{pub}}_{\varepsilon '}(f)\), this inequality holding in \(\mathbb {N}_\infty \). That is, permitting a larger error tolerance can only decrease the worst-case number of bits that must be exchanged.
Suppose \(X\times Y\) is a finite probability space, whose probability measure we write \(\mathbb {P}\), let \(m\in \mathbb {N}\), and let \(p\) be a public-coin protocol whose shared randomness is a coin tape in \(\mathrm{CoinTape}(m)\), with Alice’s private input drawn from \(X\), Bob’s from \(Y\), and outputs in \(\alpha \). Write \(\mathrm{vol}\) for the measure on \(\mathrm{CoinTape}(m)\), and for a coin tape \(\omega \) and inputs \(x,y\) let \(\mathrm{rrun}(p,x,y,\omega )\) be the randomized execution of \(p\). Then for every function \(f\colon X\times Y\to \alpha \),
On the left one first measures, for each fixed coin tape \(\omega \), the \(\mathbb {P}\)-probability of the inputs on which \(p\) disagrees with \(f\), and then averages over \(\omega \); on the right one first measures, for each fixed input pair \((x,y)\), the set of coin tapes on which \(p\) disagrees with \(f\), and then averages over the inputs.
A public-coin finite-message protocol over shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \) is a deterministic finite-message protocol in which Alice’s effective input is \(\Omega \times X\) and Bob’s effective input is \(\Omega \times Y\).
Given a message function \(f : X \to \Omega \to \beta \) and a continuation \(P : \beta \to \mathrm{Protocol}\), this constructs the protocol node in which Alice sends a \(\beta \)-valued message determined by her private input \(x\) and the shared randomness \(\omega \), then continues with \(P\).
Let \(\Omega \) be a probability space, and let \(p\) be a public-coin finite-message protocol with shared randomness in \(\Omega \), private inputs from \(X\) and \(Y\), and outputs in \(\alpha \). For every function \(f : X \to Y \to \alpha \) and every real number \(\varepsilon \), the assertion that \(p\) \(\varepsilon \)-computes \(f\) coincides with the assertion that \(p\) \(\varepsilon \)-satisfies the predicate \((x, y, a) \mapsto (a = f\, x\, y)\); that is, the two propositions are equal.
A protocol \(p\) \(\varepsilon \)-satisfies a predicate \(Q : X \to Y \to \alpha \to \mathrm{Prop}\) if for every input pair \((x, y)\),
Given a protocol \(p : \mathrm{Protocol}\; \Omega \; X\; Y\; \alpha \) and a family \(q : \alpha \to \mathrm{Protocol}\; \Omega \; X\; Y\; \beta \), the protocol \(p.\mathrm{bind}\; q\) replaces each output \(a\) of \(p\) with the protocol \(q\; a\), both sharing the same randomness \(\Omega \).
Let \(p\) be a public-coin finite-message protocol with shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \), and let \(q\) be a family of such protocols indexed by \(\alpha \) with output type \(\beta \), all sharing the same randomness space \(\Omega \). Then for every private input \(x \in X\), every private input \(y \in Y\), and every shared random string \(\omega \in \Omega \), the randomized execution of the sequential composition \(p.\mathrm{bind}\, q\) agrees with first executing \(p\) and then executing the selected continuation:
Given a message function \(f : Y \to \Omega \to \beta \) and a continuation \(P : \beta \to \mathrm{Protocol}\), this constructs the protocol node in which Bob sends a \(\beta \)-valued message determined by his private input \(y\) and the shared randomness \(\omega \), then continues with \(P\).
Given \(h : \Omega ' \to \Omega \) and a protocol \(p : \mathrm{Protocol}\; \Omega \; X\; Y\; \alpha \), the protocol \(\mathrm{comapRandomness}\; h\; p\) is a \(\mathrm{Protocol}\; \Omega '\; X\; Y\; \alpha \) that runs \(p\) after pre-composing the shared randomness with \(h\).
Let \(p\) be a public-coin finite-message protocol with shared-randomness space \(\Omega \), private inputs \(X\) and \(Y\), and outputs in \(\alpha \), and let \(h : \Omega ' \to \Omega \) be any map. Form the protocol over the new randomness space \(\Omega '\) by replacing each player’s effective input \(\Omega \times X\) or \(\Omega \times Y\) with \(\Omega ' \times X\) or \(\Omega ' \times Y\) and precomposing every message function with \(h\) on the randomness coordinate (leaving the private inputs untouched). Then this reindexed protocol has exactly the same worst-case communication complexity as \(p\).
Let \(h : \Omega ' \to \Omega \) be a map between shared-randomness spaces, and let \(p\) be a public-coin finite-message protocol over randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \). Form the reindexed protocol over \(\Omega '\) that pre-composes the shared randomness with \(h\) before running \(p\). Then for all private inputs \(x \in X\) and \(y \in Y\) and all shared randomness \(\omega \in \Omega '\), executing the reindexed protocol on inputs \(x, y\) with randomness \(\omega \) produces the same output as executing \(p\) on inputs \(x, y\) with randomness \(h(\omega )\).
Given finite types \(X\) and \(Y\) and real parameters \(\varepsilon , c\), the number of random samples required for derandomization via Chernoff and union bound is
where \(\lceil \cdot \rceil _{+}\) denotes the natural-number ceiling.
Let \(\Omega \) be a finite probability space and let \(X\) and \(Y\) be finite input sets. Let \(p\) be a public-coin finite-message protocol over shared randomness \(\Omega \) with private inputs in \(X\) and \(Y\) and output type \(\alpha \), let \(f : X \to Y \to \alpha \), and let \(\varepsilon , c \in \mathbb {R}\) with \(c \gt 1\). Suppose \(p\) \(\varepsilon \)-computes \(f\), so that for every input pair \((x,y)\) the probability over \(\omega \) that the randomized execution \(\mathrm{rrun}(p,x,y,\omega )\) differs from \(f(x,y)\) is at most \(\varepsilon \). Then, writing \(t = t(X,Y,\varepsilon ,c)\) for the derandomization sample count, there exists a family \(\omega _0, \dots , \omega _{t-1} \in \Omega \) of shared-randomness values such that for every \((x,y)\),
Given a function \(g : \alpha \to \beta \) and a public-coin protocol \(p : \mathrm{Protocol}\; \Omega \; X\; Y\; \alpha \), the protocol \(p.\mathrm{map}\; g\) is a \(\mathrm{Protocol}\; \Omega \; X\; Y\; \beta \) that applies \(g\) to the output of \(p\) after every run, without changing the communication structure.
Let \(p\) be a public-coin finite-message protocol over a shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \), and let \(g : \alpha \to \beta \) be any function. Then the protocol \(p.\mathrm{map}\, g\), which runs \(p\) and applies \(g\) to its output, has the same communication complexity as \(p\):
Let \(p\) be a public-coin finite-message protocol over shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \), and let \(g\colon \alpha \to \beta \) be any function. Write \(\mathrm{map}(g,p)\) for the protocol with the same communication structure as \(p\) that applies \(g\) to whatever value \(p\) outputs, and write \(\mathrm{rrun}(q,x,y,\omega )\) for the result of executing a protocol \(q\) on private inputs \(x\) and \(y\) with shared randomness \(\omega \). Then for all \(x\in X\), \(y\in Y\), and \(\omega \in \Omega \),
Given a public-coin finite-message protocol \(p\) over a probability space \(\Omega \) that \(\varepsilon \)-computes \(f : X \to Y \to \alpha \), and a constant \(c \gt 1\), the Newman protocol is a private-coin finite-message protocol in which Alice draws a random index \(i\) from \(\mathtt{newmanIndexSpace}\; X\; Y\; \varepsilon \; c\) and sends it to Bob; both players then run the \(i\)-th seed from a fixed table of good randomness values (obtained via a Chernoff and union-bound argument) through the original public-coin protocol deterministically.
Let \(\Omega \) be a finite probability space, let \(X\) and \(Y\) be finite types, and let \(p\) be a public-coin finite-message protocol over shared randomness \(\Omega \) with inputs in \(X\) and \(Y\) and outputs in a type \(\alpha \). Fix a target function \(f : X \times Y \to \alpha \), a real error bound \(\varepsilon \), and a real constant \(c \gt 1\), and suppose that \(p\) \(\varepsilon \)-computes \(f\), meaning that for every input pair \((x,y)\) the probability over the shared randomness that \(p\) outputs something other than \(f(x,y)\) is at most \(\varepsilon \). Then the associated Newman protocol — the private-coin protocol in which Alice draws an index into the fixed table of good random seeds, sends it to Bob, and both run \(p\) deterministically on that seed — \((c\varepsilon )\)-computes \(f\): for every input pair \((x,y)\) the probability over the private randomness that its output differs from \(f(x,y)\) is at most \(c\varepsilon \).
Let \(X\) and \(Y\) be finite types, let \(\Omega \) be a finite probability space, and let \(p\) be a public-coin finite-message protocol over \(\Omega \) with private inputs in \(X\) and \(Y\) and outputs in a type \(\alpha \). Fix a target function \(f : X \to Y \to \alpha \), real parameters \(\varepsilon \) and \(c\) with \(c \gt 1\), and suppose that \(p\) \(\varepsilon \)-computes \(f\). Then the communication complexity of the associated Newman protocol equals
where \(t(X, Y, \varepsilon , c)\) is the derandomization sample count and \(\mathrm{complexity}(p)\) is the communication complexity of \(p\); that is, running the Newman protocol costs exactly the cost of \(p\) together with the number of bits Alice needs to name one seed in the table of size \(t(X, Y, \varepsilon , c)\).
Let \(\Omega \) be a shared-randomness space, let \(X\) and \(Y\) be the private input sets of Alice and Bob, let \(\alpha \) be an output type, and let \(p\) be a binary public-coin protocol over shared randomness \(\Omega \) with these private inputs and this output — that is, a deterministic protocol whose players branch on Boolean messages, with Alice reading \(\Omega \times X\) and Bob reading \(\Omega \times Y\). Embedding \(p\) into the finite-message framework by regarding each Boolean message as an element of the two-element type \(\mathrm{Bool}\) yields a finite-message protocol whose worst-case communication complexity equals the worst-case communication complexity of \(p\).
For every public-coin protocol \(p\) over shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \), there exists a public-coin finite-message protocol \(P\) over the same data that computes exactly the same output under randomized execution and has the same communication complexity. Precisely, \(\mathrm{rrun}(P, x, y, \omega ) = \mathrm{rrun}(p, x, y, \omega )\) for all private inputs \(x \in X\), \(y \in Y\) and every shared random string \(\omega \in \Omega \), and the communication complexity of \(P\) equals that of \(p\).
Let \(\Omega \) be a shared-randomness space, let \(X\) and \(Y\) be the private input sets of the two players, and let \(\alpha \) be an output type. Let \(p\) be a binary public-coin protocol over these types, that is, a deterministic protocol whose players receive inputs in \(\Omega \times X\) and \(\Omega \times Y\). Embedding \(p\) into the finite-message framework—by treating each Boolean message as a message drawn from the two-element type—and then running the result yields the same output as running \(p\) directly: for every private input \(x \in X\), every private input \(y \in Y\), and every shared random string \(\omega \in \Omega \), the finite-message embedding of \(p\), executed on the paired inputs \((\omega , x)\) and \((\omega , y)\), returns the same value in \(\alpha \) as \(p\) executed on those same paired inputs.
Given a family of protocols \(p : (i : \mathrm{Fin}\; k) \to \mathrm{Protocol}\; (\Omega _i)\; X\; Y\; (\alpha _i)\) with heterogeneous randomness and output types, the \(\pi \)-protocol \(\mathrm{pi}\; p\) is a \(\mathrm{Protocol}\; (\prod _i \Omega _i)\; X\; Y\; (\prod _i \alpha _i)\) that runs each \(p_i\) independently using its own component of the shared randomness tuple.
Let \(k\) be a natural number, and for each index \(i \in \mathrm{Fin}\, k\) let \(p_i\) be a public-coin finite-message protocol with shared-randomness space \(\Omega _i\), common private input sets \(X\) and \(Y\), and output type \(\alpha _i\). Then the communication complexity of their \(k\)-fold product — the protocol \(\mathrm{pi}\, p\) over the combined randomness space \(\prod _i \Omega _i\) that runs each \(p_i\) on its own randomness component and returns the tuple of outputs — is the sum of the complexities of the factors:
Fix \(k \in \mathbb {N}\), shared private-input sets \(X\) and \(Y\), and, for each index \(i \in \mathrm{Fin}\, k\), a randomness space \(\Omega _i\) and an output type \(\alpha _i\). Let \(p = (p_i)_{i}\) be a family of public-coin finite-message protocols, where \(p_i\) has randomness space \(\Omega _i\), private inputs from \(X\) and \(Y\), and outputs in \(\alpha _i\), and let \(\mathrm{pi}\, p\) be their \(k\)-fold product, a public-coin protocol with randomness space \(\prod _i \Omega _i\) and output type \(\prod _i \alpha _i\). Then for all inputs \(x \in X\), \(y \in Y\) and every randomness tuple \(\omega = (\omega _i)_i \in \prod _i \Omega _i\), the randomized execution of the product on \((x, y, \omega )\) is the tuple of the individual executions:
Let \((\Omega ', \mu )\) be a probability space, let \(t \ge 1\), and let \(Y_0, \dots , Y_{t-1} : \Omega ' \to \mathbb {R}\) be mutually independent, almost everywhere measurable random variables that almost surely take values in \([0,1]\). Suppose \(\mathbb {E}[Y_i] \le \varepsilon \) for every \(i\), where \(0 \le \varepsilon \), and let \(c \gt 1\). Then
Given protocols \(p_1 : \mathrm{Protocol}\; \Omega _1\; X\; Y\; \alpha _1\) and \(p_2 : \mathrm{Protocol}\; \Omega _2\; X\; Y\; \alpha _2\), the product protocol \(\mathrm{prod}\; p_1\; p_2\) is a \(\mathrm{Protocol}\; (\Omega _1 \times \Omega _2)\; X\; Y\; (\alpha _1 \times \alpha _2)\) that runs \(p_1\) and \(p_2\) independently on the same inputs, using the first and second components of the shared randomness pair respectively.
Let \(p_1\) and \(p_2\) be public-coin finite-message protocols over the common private input sets \(X\) and \(Y\), with shared-randomness spaces \(\Omega _1\) and \(\Omega _2\) and output types \(\alpha _1\) and \(\alpha _2\) respectively. Then the communication complexity of their product protocol \(\mathrm{prod}\; p_1\; p_2\), which runs \(p_1\) and \(p_2\) independently on the same inputs, equals the sum of the individual complexities:
Let \(p_1\) be a public-coin finite-message protocol with shared-randomness space \(\Omega _1\) and output type \(\alpha _1\), and let \(p_2\) be one with shared-randomness space \(\Omega _2\) and output type \(\alpha _2\), both over the same private input sets \(X\) (Alice’s) and \(Y\) (Bob’s). Then for every pair of inputs \(x \in X\), \(y \in Y\) and every joint random string \(\omega = (\omega _1, \omega _2) \in \Omega _1 \times \Omega _2\), the randomized execution of the product protocol \(\mathrm{prod}(p_1, p_2)\) decomposes into the executions of its factors:
that is, each component is obtained by running the corresponding factor on the same inputs \(x\), \(y\) with the matching component of the shared randomness.
Given a public-coin protocol \(p_1 : \mathrm{Protocol}\; \Omega \; X\; Y\; \alpha _1\) and a deterministic protocol \(p_2 : \mathrm{Deterministic.FiniteMessage.Protocol}\; X\; Y\; \alpha _2\), the protocol \(\mathrm{prodDet}\; p_1\; p_2\) is a \(\mathrm{Protocol}\; \Omega \; X\; Y\; (\alpha _1 \times \alpha _2)\) that runs both protocols on the same inputs and pairs their outputs, reusing the same randomness \(\Omega \).
Let \(\Omega \) be a shared-randomness space and let \(X\), \(Y\) be private input sets. Let \(p_1\) be a public-coin finite-message protocol over \(\Omega \), \(X\), \(Y\) with output type \(\alpha _1\), and let \(p_2\) be a deterministic finite-message protocol over \(X\), \(Y\) with output type \(\alpha _2\). Form their product: the public-coin protocol over \(\Omega \), \(X\), \(Y\) with output type \(\alpha _1 \times \alpha _2\) that runs both \(p_1\) and \(p_2\) on the same inputs, sharing the randomness \(\Omega \), and returns the pair of their outputs. Then the worst-case communication complexity of this product protocol equals the sum of the complexities of \(p_1\) and \(p_2\).
Let \(p_1\) be a public-coin finite-message protocol over shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha _1\), and let \(p_2\) be a deterministic finite-message protocol over \(X\) and \(Y\) with output type \(\alpha _2\). Write \(p_1 \times p_2\) for their product: the public-coin protocol with output type \(\alpha _1 \times \alpha _2\) that runs both on the same private inputs and the same shared randomness and pairs their outputs. Denote by \(p_1(x, y; \omega )\) the output of executing \(p_1\) on private inputs \(x\) and \(y\) with shared randomness \(\omega \), and by \(p_2(x, y)\) the output of executing the deterministic protocol \(p_2\) on \(x\) and \(y\). Then for every pair of private inputs \(x \in X\), \(y \in Y\) and every value \(\omega \in \Omega \) of the shared randomness,
Given \(p : \mathrm{Protocol}\; \Omega \; X\; Y\; \alpha \) and \(q : \alpha \to \mathrm{Protocol}\; \Omega '\; X\; Y\; \beta \), the protocol \(\mathrm{rbind}\; p\; q\) is a \(\mathrm{Protocol}\; (\Omega \times \Omega ')\; X\; Y\; \beta \) that first runs \(p\) using the \(\Omega \)-component of the shared randomness, then runs \(q\) on the result using the independent \(\Omega '\)-component.
Let \(p\) be a public-coin finite-message protocol over shared randomness \(\Omega \), private inputs \(X\) and \(Y\), and outputs \(\alpha \), and let \(q\) assign to each output value \(a \in \alpha \) a public-coin finite-message protocol \(q(a)\) over fresh, independent shared randomness \(\Omega '\), with the same private inputs \(X\) and \(Y\) and outputs in \(\beta \). If there is a constant \(c \in \mathbb {N}\) such that every continuation \(q(a)\) has communication complexity exactly \(c\), then the protocol \(\mathrm{rbind}(p, q)\) — which runs \(p\) on the \(\Omega \)-component of the randomness and then runs the selected \(q(a)\) on the independent \(\Omega '\)-component — has communication complexity \(\mathrm{complexity}(p) + c\).
Let \(p\) be a public-coin finite-message protocol over shared-randomness space \(\Omega \) with private inputs in \(X\), \(Y\) and outputs in \(\alpha \), and let \(q\) assign to each output \(a \in \alpha \) a public-coin finite-message protocol \(q(a)\) over shared-randomness space \(\Omega '\) with the same private input sets and outputs in \(\beta \). Then for all private inputs \(x \in X\) and \(y \in Y\) and every randomness pair \(\omega = (\omega _1, \omega _2) \in \Omega \times \Omega '\), the randomized execution of the fresh-randomness bind \(\mathrm{rbind}(p, q)\) satisfies
That is, running the bound protocol with randomness \(\omega \) is the same as first running \(p\) on \(x, y\) with the randomness \(\omega _1\) to obtain an output, and then running the protocol \(q\) assigns to that output on \(x, y\) with the independent randomness \(\omega _2\).
Let \(p\) be a public-coin finite-message protocol over shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \), and let \(x \in X\), \(y \in Y\), and \(\omega \in \Omega \). Then the randomized execution of \(p\) on inputs \(x\) and \(y\) with shared randomness \(\omega \) agrees with the deterministic execution of \(p\) on Alice’s effective input \((\omega , x)\) and Bob’s effective input \((\omega , y)\):
Given a public-coin finite-message protocol \(p\) over an arbitrary finite probability space \(\Omega \) and a slack \(\delta \gt 0\), this definition produces a natural number \(n\) together with a finite-message protocol over \(\mathrm{CoinTape}(n)\) that approximates \(p\). The construction works by pulling back \(p\) along a measure-approximating map \(\varphi : \mathrm{CoinTape}(n) \to \Omega \) provided by CommunicationComplexity.Internal.single_coin_approx.
Let \(p\) be a public-coin finite-message protocol over a finite probability space \(\Omega \), with private input sets \(X\) and \(Y\) and outputs in \(\alpha \); let \(Q : X \to Y \to \alpha \to \mathrm{Prop}\) be a predicate; and let \(\varepsilon \) and \(\delta \) be real numbers with \(\delta \gt 0\). If \(p\) approximately satisfies \(Q\) with error at most \(\varepsilon \), then the coin-tape protocol obtained by converting \(p\) with slack \(\delta \) approximately satisfies \(Q\) with error at most \(\varepsilon + \delta \).
Let \(\Omega \) be a finite probability space, and let \(p\) be a public-coin finite-message protocol over shared randomness \(\Omega \) with input sets \(X\), \(Y\) and output type \(\alpha \). For every real number \(\delta \gt 0\), the coin-tape conversion of \(p\) at slack \(\delta \) yields a length \(n\) and a finite-message protocol over \(\mathrm{CoinTape}(n)\) whose communication complexity equals the communication complexity of \(p\).
Given a public-coin finite-message protocol \(p\) and a fixed randomness sample \(\omega : \Omega \), PublicCoin.FiniteMessage.Protocol.toDeterministic \(p~ \omega \) is the deterministic finite-message protocol obtained by substituting \(\omega \) for the shared randomness (via comap with \(\mathrm{Prod.mk}\, \omega \)).
Let \(p\) be a public-coin finite-message protocol over a shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \), and let \(\omega \in \Omega \) be any fixed randomness sample. Then the deterministic finite-message protocol obtained by substituting \(\omega \) for the shared randomness has the same communication complexity as \(p\) itself.
Let \(p\) be a public-coin finite-message protocol over a shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \). Fix a randomness sample \(\omega \in \Omega \) and private inputs \(x \in X\) and \(y \in Y\), and let \(p_\omega \) denote the deterministic finite-message protocol obtained from \(p\) by substituting \(\omega \) for the shared randomness. Then executing \(p_\omega \) on \(x\) and \(y\) yields the same output as the randomized execution of \(p\) on \(x\) and \(y\) with shared randomness \(\omega \):
Let \(\Omega \) be a shared-randomness space and let \(p\) be a public-coin finite-message protocol with Alice’s private input in \(X\), Bob’s in \(Y\), and output in \(\alpha \); that is, a deterministic finite-message protocol whose effective input types are \(\Omega \times X\) for Alice and \(\Omega \times Y\) for Bob, in which each message is an element of an arbitrary finite nonempty type \(\beta \) costing \(\lceil \log _2 \left\lvert \beta \right\rvert \rceil \) bits. Then the binary public-coin protocol obtained from \(p\) by encoding every \(\beta \)-valued message as \(\lceil \log _2 \left\lvert \beta \right\rvert \rceil \) bits has communication complexity exactly equal to that of \(p\).
Let \(p\) be a public-coin finite-message protocol with shared-randomness space \(\Omega \), private input sets \(X\) and \(Y\), and output type \(\alpha \), and let \(q\) be the binary public-coin protocol obtained from \(p\) by encoding each finite-alphabet message as a string of bits. Then for every pair of private inputs \(x \in X\) and \(y \in Y\) and every shared random string \(\omega \in \Omega \), the two protocols produce the same output when executed on those inputs with that randomness: the randomized execution of \(q\) on \(x\), \(y\), \(\omega \) equals the randomized execution of \(p\) on \(x\), \(y\), \(\omega \).
Let \(f : X \to Y \to \alpha \) be a function, let \(\varepsilon \in \mathbb {R}\), and let \(m \in \mathbb {N}\). Write \(\mathrm{R}^{\mathrm{pub}}_\varepsilon (f)\) for the \(\varepsilon \)-error public-coin randomized communication complexity of \(f\), namely the infimum in \(\mathbb {N}\cup \{ \infty \} \) of the worst-case bit complexity over all public-coin protocols that compute \(f\) with error at most \(\varepsilon \) on every input. Then \(m \le \mathrm{R}^{\mathrm{pub}}_\varepsilon (f)\) if and only if, for every \(n \in \mathbb {N}\) and every public-coin protocol \(p\) using a coin tape of length \(n\), whenever \(p\) \(\varepsilon \)-computes \(f\) its communication complexity satisfies \(m \le \mathrm{complexity}(p)\).
Let \(\mu \) be a finite probability space on \(X \times Y\), let \(g : X \to Y \to \mathrm{Bool}\), and let \(\varepsilon , \gamma \in \mathbb {R}\) and \(n \in \mathbb {N}\). Suppose that every rectangle \(R = A \times B \subseteq X \times Y\) satisfies \(\left\lvert \mathrm{disc}_\mu (g, R)\right\rvert \le \gamma \), where \(\mathrm{disc}_\mu (g, R) = \mathbb {E}_{(x,y)\sim \mu }\bigl[\mathbf{1}_R(x,y)\cdot \sigma (g(x,y))\bigr]\) and \(\sigma \) sends \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\). If moreover \(2^n \cdot \gamma \lt 1 - 2\varepsilon \), then \(n \lt \mathrm{R}^{\mathrm{pub}}_\varepsilon (g)\), the \(\varepsilon \)-error public-coin randomized communication complexity of \(g\).
Let \(f : X \to Y \to \alpha \) be a target function, let \(\varepsilon \in \mathbb {R}\), let \(n \in \mathbb {N}\), and let \(\mu \) be a finite probability space on \(X \times Y\). Suppose that every deterministic protocol \(p\) whose communication complexity is at most \(n\) has distributional error strictly greater than \(\varepsilon \) against \(f\) under \(\mu \); that is, the \(\mu \)-probability that \(p\) disagrees with \(f\) exceeds \(\varepsilon \). Then the \(\varepsilon \)-error public-coin randomized communication complexity of \(f\) is strictly greater than \(n\).
Let \(X\) and \(Y\) be finite types, let \(f : X \to Y \to \alpha \), and let \(\varepsilon , \varepsilon ', c \in \mathbb {R}\) satisfy \(1 \lt c\) and \(c\varepsilon \lt \varepsilon '\). Then the private-coin communication complexity of \(f\) at error tolerance \(\varepsilon '\) does not exceed the public-coin communication complexity of \(f\) at error tolerance \(\varepsilon \) increased by \(\lceil \log _2 t(X, Y, \varepsilon , c)\rceil \), where \(t(X, Y, \varepsilon , c)\) is the derandomization sample count; that is,
with both complexities taken as values in \(\mathbb {N}\cup \{ \infty \} \).
Given finite input types \(X\) and \(Y\) and parameters \(\varepsilon , c \in \mathbb {R}\), the Newman index space is \(\mathrm{Fin}(\mathtt{derandomizationSamples}\; X\; Y\; \varepsilon \; c)\), i.e. the finite type indexing the table of random seeds that Alice samples from in the Newman reduction.
The \(\varepsilon \)-error one-way public-coin communication complexity of \(f : X \to Y \to \alpha \) is the infimum (in \(\mathbb {N}_\infty \)) of the one-way message cost \(p.\mathrm{cost}\) over all \(n \in \mathbb {N}\) and all protocols \(p\) over the coin-tape space \(\mathrm{CoinTape}(n)\) that \(\varepsilon \)-compute \(f\).
Fix a function \(f : X \to Y \to \alpha \), a real error bound \(\varepsilon \), and a natural number \(m\). Then the \(\varepsilon \)-error one-way public-coin communication complexity of \(f\) is at most \(m\) (as an element of \(\mathbb {N}_\infty \)) if and only if there exist a length \(n \in \mathbb {N}\) and a one-way public-coin protocol \(p\) over the coin-tape space \(\mathrm{CoinTape}(n)\) such that \(p\) \(\varepsilon \)-computes \(f\) and the bit cost of \(p\) is at most \(m\).
Fix a function \(f : X \to Y \to \alpha \) and real error tolerances \(\varepsilon ', \varepsilon \) with \(\varepsilon ' \le \varepsilon \). Then the \(\varepsilon \)-error one-way public-coin communication complexity of \(f\) is at most its \(\varepsilon '\)-error one-way public-coin communication complexity, where each such complexity is the infimum (in \(\mathbb {N}\cup \{ \infty \} \)) of the bit cost over all coin-tape lengths \(n\) and all one-way public-coin protocols that compute \(f\) to the corresponding error tolerance. In other words, relaxing the permitted error can only decrease (or leave unchanged) the required communication.
Let \(X \times Y\) be a finite set equipped with a probability measure \(\mu \), let \(p\) be a one-way public-coin protocol with inputs in \(X\) and \(Y\), outputs in \(\alpha \), and shared randomness ranging over coin tapes of length \(m\), and let \(f : X \times Y \to \alpha \) be a target function. Write \(\mathrm{rrun}(p, x, y, \omega )\) for the output produced by running \(p\) on inputs \(x\) and \(y\) under coin tape \(\omega \), and let \(\nu \) be the measure carried by the coin tapes of length \(m\). Then
where each measure is read as a real number. That is, integrating the \(\mu \)-probability of failure over the choice of coin tape yields the same value as integrating the \(\nu \)-probability of failure over the choice of input pair.
Let \(f : X \to Y \to \alpha \) be a function, let \(\varepsilon \in \mathbb {R}\), and let \(m \in \mathbb {N}\). Then \(m\) is a lower bound for the \(\varepsilon \)-error one-way public-coin communication complexity of \(f\), in the sense that \(m \le \mathrm{CC}(f,\varepsilon )\) as elements of \(\mathbb {N}_\infty \), if and only if every protocol \(p\) over a coin tape \(\mathrm{CoinTape}(n)\) of any length \(n \in \mathbb {N}\) that \(\varepsilon \)-computes \(f\) has bit cost at least \(m\).
Let \(f : X \to Y \to \alpha \) be a function, let \(\varepsilon \in \mathbb {R}\), and let \(n \in \mathbb {N}\). Suppose there is a finite probability distribution \(\mu \) on \(X \times Y\) under which every deterministic one-way protocol \(p\) of communication cost at most \(n\) has distributional error strictly greater than \(\varepsilon \); that is, \(\Pr _{(x,y)\sim \mu }\bigl[\, p\text{ outputs a value}\neq f(x,y)\, \bigr] \gt \varepsilon \) for every such \(p\). Then the \(\varepsilon \)-error one-way public-coin communication complexity of \(f\) is strictly greater than \(n\).
A one-way public-coin protocol over shared randomness type \(\Omega \) and input types \(X\), \(Y\) with output type \(\alpha \) is defined as a deterministic one-way protocol whose Alice-input is \(\Omega \times X\) and whose Bob-input is \(\Omega \times Y\). The shared randomness \(\omega \in \Omega \) is thus treated as a component of both players’ inputs.
A protocol \(p\) \(\varepsilon \)-computes a function \(f : X \to Y \to \alpha \) if for every input pair \((x, y)\) the measure of the set of shared strings on which \(p\) errs is at most \(\varepsilon \):
Given a protocol \(p\), inputs \(x \in X\), \(y \in Y\), and a shared random string \(\omega \in \Omega \), the execution \(\mathrm{rrun}(p, x, y, \omega )\) runs \(p\) on the joint inputs \((\omega , x)\) for Alice and \((\omega , y)\) for Bob, returning the decoded output \(p.\mathrm{decode}(p.\mathrm{send}(\omega , x),\, (\omega , y))\).
Given a one-way public-coin protocol \(p\) over a randomness space \(\Omega \) and a fixed coin outcome \(\omega \in \Omega \), toDeterministic \(p\) \(\omega \) is the deterministic one-way protocol obtained by substituting \(\omega \) for the public randomness: it uses the same message type and sends/decodes by fixing \(\omega \) in \(p\).
Let \(p\) be a one-way public-coin protocol over a shared randomness type \(\Omega \), with input types \(X\), \(Y\) and output type \(\alpha \), and let \(\omega \in \Omega \) be a fixed coin outcome. Then the deterministic one-way protocol obtained from \(p\) by substituting \(\omega \) for the public randomness has the same communication cost as \(p\); that is, both have cost \(\lceil \log _2 |M| \rceil \) bits, where \(M\) is the shared message codebook.
Let \(p\) be a one-way public-coin protocol with shared-randomness type \(\Omega \), input types \(X\) and \(Y\), and output type \(\alpha \); let \(\omega \in \Omega \) be a coin outcome and \(x \in X\), \(y \in Y\) be inputs. Fixing the public coin at \(\omega \) turns \(p\) into a deterministic one-way protocol, and executing that deterministic protocol on inputs \(x\) and \(y\) yields the same output as running \(p\) on the joint inputs \((\omega , x)\) for Alice and \((\omega , y)\) for Bob. In both cases the output is Bob’s decoder applied to his input \((\omega , y)\) together with the message that Alice’s encoder assigns to \((\omega , x)\).
A public-coin protocol over shared randomness \(\Omega \), with Alice’s private input in \(X\) and Bob’s in \(Y\) producing an output in \(\alpha \), is a deterministic protocol whose combined input types are \(\Omega \times X\) (for Alice) and \(\Omega \times Y\) (for Bob). In other words, both players observe the same random string \(\omega \in \Omega \) alongside their private inputs.
A public-coin protocol \(p\) \(\varepsilon \)-computes a function \(f : X \to Y \to \alpha \) if for every input pair \((x, y)\),
Here the volume is taken with respect to the measure on \(\Omega \) given by the MeasureSpace instance.
Let \(p\) be a public-coin protocol over a shared-randomness space \(\Omega \) carrying a measure, with Alice’s inputs in \(X\), Bob’s in \(Y\), and outputs in \(\alpha \), let \(f : X \to Y \to \alpha \) be a function, and let \(\varepsilon \in \mathbb {R}\). Then the assertion that \(p\) \(\varepsilon \)-computes \(f\) coincides with the assertion that \(p\) \(\varepsilon \)-satisfies the predicate \(Q(x, y, a)\) given by \(a = f(x, y)\); that is, the two propositions are equal.
A public-coin protocol \(p\) \(\varepsilon \)-satisfies a predicate \(Q : X \to Y \to \alpha \to \mathrm{Prop}\) if for every input pair \((x, y)\),
Given a public-coin protocol \(p\), private inputs \(x \in X\), \(y \in Y\), and a shared random string \(\omega \in \Omega \), \(\texttt{rrun}\, p\, x\, y\, \omega \) executes the underlying deterministic protocol on the paired inputs \((\omega , x)\) and \((\omega , y)\) and returns the resulting output.
Let \(p\) be a public-coin protocol with shared randomness in \(\Omega \), Alice’s private input in \(X\), and Bob’s in \(Y\), producing an output in \(\alpha \). Then for every \(x \in X\), \(y \in Y\), and \(\omega \in \Omega \), the randomized execution of \(p\) on \(x\), \(y\) with random string \(\omega \) equals the execution of the underlying deterministic protocol on the paired inputs \((\omega , x)\) and \((\omega , y)\).
Given a public-coin protocol \(p\) over randomness space \(\Omega \) and a fixed outcome \(\omega : \Omega \), PublicCoin.Protocol.toDeterministic \(p~ \omega \) is the deterministic protocol obtained by substituting \(\omega \) for the shared randomness (implemented via comap with \(\mathrm{Prod.mk}\, \omega \)).
Let \(p\) be a public-coin communication protocol over a shared randomness space \(\Omega \), with Alice’s private input in \(X\), Bob’s in \(Y\), and output in \(\alpha \), and let \(\omega \in \Omega \) be any outcome of the shared randomness. Then the communication complexity of the deterministic protocol obtained from \(p\) by fixing the shared randomness to \(\omega \) equals the communication complexity of \(p\) itself.
Let \(p\) be a public-coin protocol over a randomness space \(\Omega \), with Alice’s private input in \(X\), Bob’s in \(Y\), and output in \(\alpha \), and fix a random outcome \(\omega \in \Omega \). For all private inputs \(x \in X\) and \(y \in Y\), executing the deterministic protocol obtained from \(p\) by fixing the shared randomness to \(\omega \) on the pair \((x, y)\) returns the same output in \(\alpha \) as the randomised execution of \(p\) on \(x\) and \(y\) with shared random string \(\omega \).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and write \(\mathrm{TV}(\mu ,\nu )\) for their total variation distance. Suppose \(\mu \) is absolutely continuous with respect to \(\nu \), and that the log-likelihood ratio \(x \mapsto \log \frac{d\mu }{d\nu }(x)\) is \(\mu \)-integrable. Then
the right-hand side being the Kullback–Leibler divergence of \(\mu \) from \(\nu \).
Let \(X\), \(Y\), and \(\alpha \) be sets and let \(g : X \times Y \to \alpha \). If \(S \subseteq X \times Y\) is a fooling set for \(g\) and \(\mathcal{P}\) is a monochromatic rectangle partition for \(g\), then the cardinality of \(S\) is at most the cardinality of \(\mathcal{P}\), that is \(\left\lvert S\right\rvert \le \left\lvert \mathcal{P}\right\rvert \), where cardinalities are compared as extended natural numbers (allowing the value \(\infty \)).
Let \(g : X \times Y \to \alpha \), let \(S \subseteq X \times Y\) be a fooling set for \(g\), and let \(\mathcal{P}\) be a monochromatic rectangle partition for \(g\). If \(\mathcal{P}\) is finite, then the number of points of \(S\) is at most the number of members of \(\mathcal{P}\); that is, \(\left\lvert S\right\rvert \le \left\lvert \mathcal{P}\right\rvert \).
A set \(S \subseteq X \times Y\) is a fooling set for \(g : X \to Y \to \alpha \) if every monochromatic rectangle with respect to \(g\) contains at most one point of \(S\), i.e., \(S \cap R\) is a subsingleton for every rectangle \(R\) that is monochromatic for \(g\).
A collection \(\mathcal{P}\) of subsets of \(X \times Y\) is a monochromatic rectangle partition for \(g : X \to Y \to \alpha \) if every member of \(\mathcal{P}\) is a rectangle, every member is monochromatic for \(g\), the members cover all of \(X \times Y\), and distinct members are disjoint.
A subset \(R \subseteq X \times Y\) is a rectangle — that is, \(R = A \times B\) for some \(A \subseteq X\) and \(B \subseteq Y\) — if and only if it is closed under crossing: for all \(x, x' \in X\) and \(y, y' \in Y\), whenever \((x, y) \in R\) and \((x', y') \in R\), both mixed pairs \((x', y)\) and \((x, y')\) also lie in \(R\).
Let \(\mathcal{P}\) be a monochromatic rectangle partition for a function \(g : X \to Y \to \alpha \), and let \(R \in \mathcal{P}\). If \((x,y)\) and \((x',y')\) both belong to \(R\), then the two "crossed" pairs \((x',y)\) and \((x,y')\) also belong to \(R\).
For two probability measures \(\mu \) and \(\nu \) on a measurable space \(\Omega \), the Radon-Nikodym density \(\mathrm{rnDensity}(\mu ,\nu )(x) \in \mathbb {R}\) is the real part of the Radon-Nikodym derivative \(\frac{d\mu }{d\nu }(x)\), obtained by converting the extended non-negative real value \(\mu .\mathrm{rnDeriv}\, \nu \, x\) to a real number via ENNReal.toReal.
Let \(\Omega \) be a finite measurable space carrying two finite measures \(\mu \) and \(\nu \) with \(\mu \ll \nu \), and let \(\omega \in \Omega \) be a point whose singleton has positive mass, \(\nu (\{ \omega \} ) \neq 0\). Then the Radon–Nikodym derivative of \(\mu \) with respect to \(\nu \) takes at \(\omega \) the value
where both sides are read as nonnegative real numbers.
Let \(\Omega \) be a finite discrete measurable space, let \(\mu \) and \(\nu \) be finite measures on \(\Omega \) with \(\mu \) absolutely continuous with respect to \(\nu \), and let \(\omega \in \Omega \). Writing \(\mu (\{ \omega \} )\) and \(\nu (\{ \omega \} )\) for the (real-valued) masses of the singleton \(\{ \omega \} \) and \(\log \frac{d\mu }{d\nu }\) for the log-likelihood ratio of \(\mu \) against \(\nu \), one has
Let \(\mu \) be a finite measure on a measurable space \(\Omega \), and let \(g \colon \Omega \to \mathbb {R}\) be measurable and integrable with mean zero, \(\int _\Omega g\, d\mu = 0\). Then the supremum of \(\left\lvert \int _S g\, d\mu \right\rvert \) taken over all measurable sets \(S \subseteq \Omega \) equals the integral of \(g\) over the region where it is nonnegative:
Let \(u,v\) be words of length \(n\) over a finite alphabet \(\alpha \), that is, functions \(\mathrm{Fin}\, n \to \alpha \). Then the Hamming distance between them equals the number of coordinates on which they differ:
Fix a word \(u\) of length \(n\) over a finite alphabet, and let \(S\) and \(T\) be two distinct sets of coordinate positions in \(\{ 0,\dots ,n-1\} \). Then the support fiber of \(u\) over \(S\) and the support fiber of \(u\) over \(T\) are disjoint: no word \(v\) can differ from \(u\) at exactly the positions in \(S\) while also differing from \(u\) at exactly the positions in \(T\).
Let \(t\) be a binary tree with more than one leaf, and write \(\left\lvert u\right\rvert \) for the number of leaves of a tree \(u\). Then \(t\) has a subtree \(s\) whose leaf count satisfies
Equivalently, \(s\) contains at least one third and strictly fewer than two thirds of the leaves of \(t\).
Let \(t\) be a binary tree, and write \(\ell (t)\) for its number of leaves. Let \(n\) be a natural number with \(n \gt 1\), and suppose that \(3\, \ell (t) \ge 2n\). Then \(t\) has a subtree \(s\) whose number of leaves satisfies
TreeIsSubtree s t is an inductively defined proposition asserting that \(s\) is a subtree of \(t\). It has two constructors: reflexivity (\(s\) is a subtree of itself) and a descent rule (\(s\) is a subtree of the left child \(\ell \) of \(\texttt{node}(v,\ell ,r)\), hence a subtree of the whole node; an analogous right rule also exists in the implementation).
The total variation distance between two probability measures \(\mu \) and \(\nu \) on \(\Omega \) is defined as \(\tfrac {1}{2}\) times the total variation norm of the signed measure \(\mu - \nu \), i.e. \(\mathrm{TV}(\mu ,\nu ) = \tfrac {1}{2}\| \mu - \nu \| _{\mathrm{TV}}\).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \). Then for every measurable set \(S \subseteq \Omega \),
where \(\mathrm{TV}(\mu ,\nu ) = \tfrac {1}{2}\lVert \mu - \nu \rVert _{\mathrm{TV}}\) is the total variation distance between \(\mu \) and \(\nu \).
Let \(\alpha \) be a finite type and let \(a : \alpha \to \mathbb {R}\) be a real-valued function whose entries sum to zero, \(\sum _{x \in \alpha } a(x) = 0\). Then for every subset \(S \subseteq \alpha \),
Let \(\alpha \) be a finite type and let \(a : \alpha \to \mathbb {R}\) be a family of real numbers whose total sum vanishes, \(\sum _{x} a(x) = 0\). Then
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\mu - \nu \) denote their signed measure difference. Then for every measurable set \(S \subseteq \Omega \),
where \(\left\lvert \mu - \nu \right\rvert \) is the total variation measure of \(\mu - \nu \).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\mu - \nu \) denote their signed measure difference. Then for every measurable set \(S \subseteq \Omega \),
where \(\left\lvert \mu - \nu \right\rvert \) is the total variation measure of \(\mu - \nu \) and \(\left\lvert \mu - \nu \right\rvert (\Omega )\) is its total mass.
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\mu - \nu \) denote their signed measure difference, with total variation measure \(\left\lvert \mu -\nu \right\rvert \). Then there exists a measurable set \(S \subseteq \Omega \) such that
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\mu - \nu \) denote their signed measure difference. Then there exists a measurable set \(S \subseteq \Omega \) at which
where \(\left\lvert \mu - \nu \right\rvert \) is the total variation measure of \(\mu - \nu \); that is, some event realizes exactly half the total variation of \(\mu - \nu \) over the whole space.
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\mu -\nu \) denote their signed measure difference. Writing \(\left\lvert \mu -\nu \right\rvert \) for its total variation measure and \((\mu -\nu )^{+}\) for the positive part of its Jordan decomposition, one has
where both sides are evaluated as real numbers on the whole space \(\Omega \).
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and let \(\mu - \nu \) be their signed measure difference. Writing \((\mu -\nu )^{+}\) and \((\mu -\nu )^{-}\) for the positive and negative parts of the Jordan decomposition of \(\mu - \nu \), the two parts assign the same total real mass to the whole space:
Let \(\mu \) be a probability measure on the two-point space \(\mathrm{Bool} = \{ \mathrm{true}, \mathrm{false}\} \). Then the mass at \(\mathrm{false}\) is one minus the mass at \(\mathrm{true}\):
Let \(\Omega \) be a finite discrete measurable space, and let \(\mu \) and \(\nu \) be probability measures on \(\Omega \). Then for every subset \(S \subseteq \Omega \),
where the sum ranges over all points of \(\Omega \) and each term contributes the singleton mass difference \(\mu (\{ \omega \} ) - \nu (\{ \omega \} )\) exactly when \(\omega \in S\) and \(0\) otherwise.
Let \(\mu _1\) be a measure on a measurable space \(\alpha \) and \(\mu _2\) a \(\sigma \)-finite measure on a measurable space \(\beta \). Then for every \(x \in \alpha \) and \(y \in \beta \), the real-valued mass that the product measure \(\mu _1 \otimes \mu _2\) assigns to the singleton \(\{ (x,y)\} \) is the product of the singleton masses:
Let \(s\) be a signed measure on a measurable space \(\Omega \), and let \(s = s^+ - s^-\) be its Jordan decomposition into a pair of (finite, mutually singular) positive measures \(s^+\) and \(s^-\). Then for every measurable set \(S \subseteq \Omega \),
where \(s^+(S)\) and \(s^-(S)\) denote the real-valued measures of \(S\) under the positive and negative parts.
Let \(a\colon \alpha \to \mathbb {R}\) be a real-valued function on a finite type \(\alpha \) whose values sum to zero, that is, \(\sum _{x\in \alpha } a(x) = 0\). Then the sum of the positive parts of \(a\) equals half the sum of its absolute values:
Let \(\alpha \) and \(\beta \) be finite types, and let \(a, c : \alpha \to \mathbb {R}\) and \(b, d : \beta \to \mathbb {R}\) be real-valued weight functions. Suppose \(b\) and \(c\) are nonnegative and each sums to one, i.e. \(b(y) \ge 0\) for all \(y\), \(c(x) \ge 0\) for all \(x\), \(\sum _{y \in \beta } b(y) = 1\) and \(\sum _{x \in \alpha } c(x) = 1\). Then the \(\ell ^1\) gap between the two product weight functions is at most the sum of the marginal \(\ell ^1\) gaps:
Let \(\Omega \) be a finite discrete measurable space, and let \(\mu \) and \(\nu \) be probability measures on \(\Omega \). Then the singleton mass differences sum to zero:
Let \(s\) be a signed measure on a measurable space \(\Omega \), with Jordan decomposition \(s = s^+ - s^-\) into its positive and negative parts \(s^+\) and \(s^-\). Then the total mass of the total variation measure \(\left\lvert s\right\rvert \) of \(s\), taken as a real number, is the sum of the total masses of the two Jordan parts:
Let \(\mu \) and \(\nu \) be probability measures on the two-point space \(\mathrm{Bool} = \{ \mathrm{true}, \mathrm{false}\} \). Then their total variation distance is the absolute difference of the masses they assign to the single point \(\mathrm{true}\):
Let \(\Omega \) be a finite discrete measurable space, and let \(\mu \) and \(\nu \) be probability measures on \(\Omega \). Then the total variation distance between \(\mu \) and \(\nu \) equals half the sum, over all points \(\omega \in \Omega \), of the absolute differences of their point masses:
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \). Then the total variation distance \(\mathrm{TV}(\mu ,\nu ) = \tfrac {1}{2}\lVert \mu -\nu \rVert _{\mathrm{TV}}\), defined as half the total variation norm of the signed measure \(\mu -\nu \), agrees with its supremum form over events,
Let \(\alpha \) and \(\beta \) be finite discrete measurable spaces, let \(\mu _1,\nu _1\) be probability measures on \(\alpha \), and let \(\mu _2,\nu _2\) be probability measures on \(\beta \). Writing \(\mathrm{TV}\) for the total variation distance and \(\mu _1\otimes \mu _2\) for the product measure, the total variation distance between the two product measures satisfies
Let \(\Omega \) be a finite set equipped with the discrete measurable structure, and let \(\mu \) and \(\nu \) be probability measures on \(\Omega \). Then the supremum form of total variation distance, \(\mathrm{TV}_{\sup }(\mu ,\nu ) = \sup \{ |\mu (S) - \nu (S)| : S \subseteq \Omega \} \), is exactly half the \(\ell ^1\) distance between the two measures across the individual points:
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and suppose that \(\mu \) is absolutely continuous with respect to \(\nu \). Write \(\frac{d\mu }{d\nu }\) for the (real-valued) Radon–Nikodym derivative of \(\mu \) with respect to \(\nu \). Then the total variation distance between \(\mu \) and \(\nu \) equals the excess mass of \(\mu \) over \(\nu \) on the region where \(\mu \) locally dominates \(\nu \):
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and suppose that \(\mu \) is absolutely continuous with respect to \(\nu \). Then the total variation distance between \(\mu \) and \(\nu \) equals one half of the density absolute integral, that is,
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and suppose \(\mu \) is absolutely continuous with respect to \(\nu \). Write \(\frac{d\mu }{d\nu }\) for the Radon–Nikodym derivative of \(\mu \) with respect to \(\nu \). Then the total variation distance, taken as the supremum
equals the excess mass of \(\mu \) over \(\nu \) on the region where \(\mu \) locally dominates:
Let \(\mu \) and \(\nu \) be probability measures on a measurable space \(\Omega \), and suppose that the Kullback–Leibler divergence \(\mathrm{KL}(\mu \, \| \, \nu )\) is finite. Then the total variation distance \(\mathrm{TV}(\mu ,\nu ) = \tfrac {1}{2}\| \mu - \nu \| _{\mathrm{TV}}\) satisfies
Let \(\Omega \) be a nonempty finite set, equipped with the uniform probability measure that assigns each point mass \(1/\left\lvert \Omega \right\rvert \), and let \(S \subseteq \Omega \). Then the measure of \(S\), as a real number, is
the number of elements of \(\Omega \) lying in \(S\) divided by the total number of elements of \(\Omega \).
Let \(\Omega \) be a nonempty finite set, equipped with its discrete measurable structure, and let \(S \subseteq \Omega \). Under the uniform probability measure on \(\Omega \), the measure of \(S\), regarded as a real number, equals
where \(\left\lvert S\right\rvert \) is the number of elements of \(\Omega \) lying in \(S\) and \(\left\lvert \Omega \right\rvert \) is the total number of elements of \(\Omega \).
Let \(\mu \) be a probability measure on a measurable space \(\Omega \), let \(f, X : \Omega \to \mathbb {R}\), and let \(t \in \mathbb {R}\). Suppose that \(f \ge 0\) pointwise, that \(\int _\Omega f\, d\mu = 1\), and that \(f\), the map \(x \mapsto f(x)\log f(x) - f(x) + 1\), the map \(x \mapsto e^{tX(x)}\), and the product \(fX\) are all \(\mu \)-integrable. Writing \(\Lambda (t) = \log \int _\Omega e^{tX}\, d\mu \) for the cumulant generating function of \(X\) under \(\mu \), one has
Let \(L\) be a language over the binary alphabet such that \(x \in L\) exactly when some binary string \(u\) with \(|u| = C \cdot 2^{(|x|+1)^c}\) has \(x\, u \in V\), and let \(M\) be a deterministic machine deciding the verifier language \(V\) within \(A(m+1)^d\). Then there are constants \(K, r\) and a nondeterministic machine deciding \(L\) within \(K\, (n + C\cdot 2^{(n+1)^c} + 1)^r\).
For every input \(x\) and every index \(i \lt |x|\), one step of the countdown emitter from the copy configuration holding the first \(i\) bits of \(x\) appends the bit at index \(i\) to the work tape and advances both heads, emitting nothing.
For every input, head position and counter word \(u\) of value \(n\), and every already-emitted word: within \((n+1)(2|u|+3)\) steps from the borrow state at cell zero, the countdown emitter halts with output equal to the already-emitted word followed by exactly \(n\) ones.
For all natural numbers \(C\) and \(c\) there exist a deterministic machine \(S\) and a constant \(B\) such that on every input \(x\), \(S\) outputs the all-ones word of length \(C \cdot 2^{(|x|+1)^c}\) within \(B\, (|x| + C\cdot 2^{(|x|+1)^c} + 1)^{c+2}\) steps.
For every input position, every counter word \(u\) and every \(j \le |u|\): from the startup-rewind configuration with work head at cell \(j-1\), exactly \(j+1\) steps of the countdown emitter reach the borrow state at cell zero with counter \(u\) and empty output.
Let \(L\) be a language over the binary alphabet and \(C, c\) natural numbers. If some nondeterministic machine decides \(L\) within time \(K(n + C\cdot 2^{(n+1)^c} + 1)^r\) for constants \(K, r\), then \(L\) belongs to \(\bigcup _e \mathrm{NTIME}(2^{n^e})\).
For every input \(x\), the \((2|x| + 2)\)-step run of the countdown emitter from its blank-tape initial configuration is the borrow-state configuration at cell zero with the full input \(x\) as counter, input head at the right boundary and empty output.
For every counter word \(u\): a full borrow-and-rewind round of the countdown emitter from cell zero costs \(2\beta + 2 \le 2|u| + 2\) steps, where \(\beta \) is the borrow position of \(u\), and ends at cell zero in the round-return state carrying the success flag, with the counter replaced by its fixed-width decrement and no output emitted.
For every counter word \(u\), every flag \(b\) and every \(j \le |u|\): from the rewind state with work head at cell \(j-1\), exactly \(j+1\) steps of the countdown emitter reach the round-return state carrying \(b\), with head at cell zero and the counter unchanged.
For every counter word \(u\) and every prefix \(p\): starting in the borrow state with work head at cell \(|p|\) over the word \(p\, u\), exactly \(\beta + 1\) steps of the countdown emitter, where \(\beta \) is the borrow position of \(u\), reach the rewind state carrying the success flag of the decrement of \(u\), with the work tape now holding \(p\) followed by the decremented word.
For every emitted prefix, every input and position, and every counter word split as \(p\, s\): one step of the countdown emitter in the borrow state with work head at cell \(|p|\) clears a leading \(1\) of \(s\) and turns to rewind with success, converts a leading \(0\) of \(s\) to \(1\) and advances, or, at the blank end of the word, turns to rewind with failure; in each case the width is preserved and nothing is emitted.
For a binary word viewed as a little-endian counter, the decrement returns a word of the same width together with a success flag: a leading \(1\) is cleared (after turning the traversed low zero bits into ones), and on underflow, when the word has value zero, the flag is false and the width is still unchanged.
Given an already-emitted output word, a physical input with an arbitrary head position, a borrow-phase control state, a work-head value and a counter word, the corresponding configuration of the countdown emitter holds the counter on its single work tape and the emitted word as physical output.
A deterministic machine with one work tape is defined that copies its binary input onto the work tape, rewinds, and then repeatedly performs a fixed-width little-endian decrement with borrow and rewind phases, emitting one \(1\) bit after each successful decrement; underflow halts silently.
Given a copy-phase flag, a physical input position, a copied word and a work-head value, the corresponding configuration of the countdown emitter is in its startup control state with the copied word on the work tape and empty output.
The value of a binary word read as a little-endian counter is defined recursively: the empty word has value \(0\), and a word with first bit \(b\) and tail \(t\) has value twice the value of \(t\), plus one exactly when \(b\) is \(1\).
If a machine \(F\) computes \(y\) from \(x\) within \(s\) steps and a machine \(G\) computes \(z\) from \(y\) within \(t\) steps, then the buffered composition of \(F\) and \(G\) computes \(z\) from \(x\) within \(s + |y| + 2 + t\) steps.
For every \(j \le |x|\): from the rewind-phase configuration of the decider wrapper with the full copy of \(x\) on tape zero and both active heads just left of cell \(j\), exactly \(j + 1\) steps reach the dispatch seam — the dispatch state over \(x\) on tape zero with all heads at the origin. This includes the empty input.
Suppose the inner machine, started in its entry state at the seam carrying \(x\) on tape zero, first reaches the exit state at a positive time \(t\) (no visit at any earlier positive time), in the canonical configuration carrying a word \(y\) on tape zero. Then the decider wrapper, from its dispatch seam, reaches the renamed exit configuration in exactly \(t\) steps; this covers the case where the entry and exit states coincide.
As long as the inner machine’s run from a configuration has not reached the designated exit state, the decider wrapper’s run from the state-renamed copy of that configuration coincides, under the renaming, with the inner machine’s run.
Suppose the inner machine, started in its entry state at the seam carrying \(x\), first returns to the exit state at a positive time \(t\) with the singleton word \([b]\) on tape zero. Then the decider wrapper computes the output \([b]\) from input \(x\) within \(2|x| + 2 + t + 1\) steps, with no earlier emission.
Let \(C\) be a finite machine with at least one work tape and designated entry and exit states, and let \(x\) be an input. After exactly \(2|x| + 2\) steps from its initial configuration on \(x\), the decider wrapper of \(C\) is in the dispatch seam configuration: the dispatch state, \(x\) on tape zero, every other work tape blank, and all heads at their starting positions.
For every word \(w\), the shift machine for \(w\) computes the function mapping each input \(x\) to \(0^{|x|} \mathbin {+\! \! +} w\) within time \((|w| + 2)(n + 1)\) on inputs of length \(n\); this time bound is monotone and linear in \(n\).
For all constants \(C, c\), the map \(n \mapsto n + C \cdot 2^{(n+1)^c}\) on the natural numbers is strictly increasing; in particular the exponential split of a padded length, when it exists, is unique. This includes coefficient zero and degree zero.
For all constants \(C, c\) there exist a finite machine and a constant \(A\) computing the split-word function (the paired prefix/suffix encoding, or the empty failure word) within time \(A (n + 1)^{c+2}\) on inputs of length \(n\).
A load-phase configuration of the decider wrapper on input \(x\) is determined by a phase index \(q \in \{ 0, 1, 2\} \), an input-head position \(p\), a word \(u\), and a head position \(h\): its control state is phase \(q\), its input head is at \(p\), tape zero holds \(u\) with its head at \(h\), every other work tape is blank with its head at zero, and the output is empty.
Let \(F\) and \(G\) be deterministic machines, and suppose \(F\) computes \(y\) from \(x\) within \(s\) steps and \(G\) computes \(z\) from \(y\) within \(t\) steps. Then the buffered composition of \(F\) and \(G\) computes \(z\) from \(x\) within \(s + |y| + 2 + t\) steps.
Suppose the comparator input has the form \(y = p\, v\) and let \(u\) be the decoded word. After \(|v| + 1\) steps from the comparison state at position \(|p|+1\), the pair comparator has halted and its output is the single bit recording whether \(u = v\).
Suppose the comparator input has the form \(y = p\, b\, b\, r\) for a bit \(b\) and words \(p, r\). From the parser state at position \(|p|+1\) with decoded word \(u\), exactly two steps of the pair comparator append \(b\) to the work tape and move to position \(|p|+3\), emitting nothing.
Suppose the comparator input has the form \(y = p\, e\) where \(e\) is the self-delimiting encoding of a pair \((u, v)\). From the parser state at position \(|p|+1\) with decoded word \(a\), exactly \(2|u|+2\) steps of the pair comparator reach the rewind state at position \(|p|+2|u|+3\) with decoded word \(a\, u\).
For every input, every decoded word \(u\), every position and every \(j \le |u|\): from the rewind state with work head at cell \(j - 1\), exactly \(j+1\) steps of the pair comparator reach the comparison state with work head at cell zero and all data unchanged.
Suppose the comparator input has the form \(y = p\, v\) and let \(u\) be the decoded word. For every \(j \le |v|\), after \(j\) steps from the comparison state at position \(|p|+1\), the comparator is at position \(|p|+j+1\) with its flag recording whether the length-\(j\) prefixes of \(u\) and \(v\) agree.
Suppose the comparator input has the form \(y = p\, 0\, 1\, v\) for words \(p, v\). From the parser state at position \(|p|+1\) with decoded word \(u\), exactly two steps of the pair comparator enter the rewind state at position \(|p|+3\), with the work tape and output untouched and the work head moved one cell left.
For all natural numbers \(C\) and \(c\) there exist a deterministic machine and a constant \(K\) computing, on every input \(x\), the pair encoding of \(x\) with the all-ones pad of length \(C \cdot 2^{(|x|+1)^c}\), within \(K\, (n + C\cdot 2^{(3(n+1))^c} + 1)^{c+2}\) steps on inputs of length \(n\).
Let \(L\) and \(V\) be languages with \(V \in \mathrm{P}\), and let \(C, c\) be natural numbers such that \(x \in L\) if and only if some \(u\) with \(|u| = C \cdot 2^{(|x|+1)^c}\) has \(x\, u \in V\). Then the exponentially padded language of \(L\) belongs to \(\mathrm{NP}\).
Let \(L\) be a language and \(C, c\) natural numbers. For every binary string \(x\), the pair encoding of \(x\) with the all-ones pad of length \(C\cdot 2^{(|x|+1)^c}\) belongs to the exponentially padded language of \(L\) if and only if \(x \in L\).
Let \(p\) be a Boolean test whose single-bit indicator is polynomial-time computable, and let \(f\) and \(g\) be polynomial-time computable functions on binary strings. Then the function that returns \(f(x)\) when \(p(x)\) holds and \(g(x)\) otherwise is polynomial-time computable.
Let \(g\) be a polynomial-time computable function on binary strings. Then the function that decodes its input as a pair \((a,b)\) and re-encodes the pair \((a, g(b))\), returning the empty word on malformed input, is polynomial-time computable.
Let \(C, c\) be natural numbers and suppose a string \(w\) decodes as a pair with first component \(x'\), which in turn decodes as a pair with first component \(x\). Then \(|x| \le |w|\), and the binary expansion of \(C \cdot 2^{(|x|+1)^c}\) has length at most \((|w|+1)^c\) plus the length of the binary expansion of \(C\).
For all natural numbers \(C, c\) and every language \(V\) in \(\mathrm{P}\), the paired verifier language for \(C\), \(c\) and \(V\), which checks both pair decodings, the exact all-ones pad of length \(C\cdot 2^{(|x|+1)^c}\), the exact witness length, and membership of the assembled word in \(V\), belongs to \(\mathrm{P}\).
A deterministic machine with one work tape is defined that, on a self-delimiting encoded pair, decodes the doubled-bit prefix onto its work tape, rewinds, then compares the stored word with the suffix bit by bit, including the final blank, and halts emitting a single bit recording whole-word equality.
For a finite machine \(M\), the nonempty-input guard of \(M\) is the machine with the same work tapes that, on empty input, emits a single rejecting (false) bit and halts, and, on nonempty input, takes one silent step and then runs \(M\) from \(M\)’s initial state.
Given natural numbers \(C, c\) and a language \(V\), the paired verifier language consists of the strings that decode as a pair \((x', u)\) whose first component decodes in turn as a pair \((x, p)\), such that \(p\) is the all-ones word of length exactly \(C\cdot 2^{(|x|+1)^c}\), the witness \(u\) has exactly that same length, and the concatenation \(x\, u\) belongs to \(V\); strings failing either decoding are excluded.
For a machine \(C\) with at least one work tape and designated entry and exit states: the wrapper machine that copies the native input onto tape zero, rewinds both heads, executes one entry action of \(C\), follows \(C\) until its first return to the exit state, and then emits the single symbol found on tape zero as its decision.
Given constants \(C, c\) and a word \(y\), the split word of \(y\) is the paired encoding of the first \(i\) symbols of \(y\) and the remaining suffix when the exponential-split search on \((C, c, |y|)\) returns \(i\), and is the empty word, signalling failure, when the search fails.
For a language \(L\) over the binary alphabet and a degree \(c\): the language of words \(y\) such that the exponential-split search with coefficient one succeeds on \(|y|\) with some prefix length \(n\) and the first \(n\) symbols of \(y\) form a word in \(L\). Words whose length admits no split are rejected; certificate bits carry no information.
Let \(a\) be a one-step action of a \(k\)-work-tape binary-alphabet Turing machine with state set \(Q\): an input-head move, a write-and-move instruction for each work tape, an optional output bit, and a successor state \(s \in Q \cup \{ \bot \} \) (\(\bot \) meaning halt). For a bit \(b\), the transformed action has the same input-head move and work-tape instructions, emits no output, and has successor state \(\mathrm{enc}(s, b')\), where \(b'\) is the bit emitted by \(a\) (or \(b\) if \(a\) emits nothing) and \(\mathrm{enc}(s, b')\) is \((s, b')\) for \(s \in Q\), halt for \(s = \bot \) and \(b' = 1\), and a dedicated loop state for \(s = \bot \) and \(b' = 0\).
Let \(M\) be a finite-state binary-alphabet Turing machine with state set \(Q\), and \(M^{\mathrm{acc}}\) its halting recognizer (states \((q,b) \in Q \times \{ 0,1\} \) plus a loop state \(\ell \), the bit recording \(M\)’s last emitted symbol). For an input \(x\) and a configuration \(C\) of \(M\) on \(x\), the configuration \(\Phi (C)\) of \(M^{\mathrm{acc}}\) on \(x\) has the same input-head position, work-tape contents and work-head positions as \(C\), an empty output tape, and state \(\mathrm{enc}(s, b)\), where \(s\) is the state of \(C\) (or \(\bot \) if halted), \(b\) is the last bit on \(C\)’s output tape (\(0\) if it is empty), and \(\mathrm{enc}(s,b)\) is \((s,b)\) for \(s \in Q\), halted for \(s=\bot , b=1\), and \(\ell \) for \(s=\bot , b=0\).
Let \(M\) be a finite-state binary-alphabet Turing machine with state set \(Q\), \(x\) an input, \(C\) a configuration of \(M\) on \(x\), and \(a\) an action of \(M\). Let \(b\) be the last bit on \(C\)’s output tape (\(0\) if empty), and let \(\Phi (C)\) be the recognizer configuration with \(C\)’s head positions and tape contents, empty output, and state \(\mathrm{enc}(s,b)\), where \(s\) is \(C\)’s state and \(\mathrm{enc}(s,b) = (s,b)\) for \(s \in Q\), halt for a halted \(s\) with \(b=1\), and a loop state for a halted \(s\) with \(b=0\). Then applying to \(\Phi (C)\) the action with \(a\)’s head moves and writes, no output, and successor state \(\mathrm{enc}(s', b')\) (\(s'\) the successor of \(a\), \(b'\) the bit \(a\) emits, or \(b\) if none) gives exactly \(\Phi (a(C))\), where \(a(C)\) is the result of applying \(a\) to \(C\).
Let \(M\) be a finite-state binary-alphabet Turing machine with state set \(Q\), and let \(M^{\mathrm{acc}}\) be its halting recognizer: same work tapes, states \((q,b) \in Q \times \{ 0,1\} \) plus a stationary loop state \(\ell \), simulating \(M\) with output suppressed and the last emitted bit recorded, halting where \(M\) halts with recorded bit \(1\) and entering \(\ell \) where it halts with bit \(0\). For a configuration \(C\) of \(M\) on an input \(x\), let \(\Phi (C)\) be the configuration of \(M^{\mathrm{acc}}\) with \(C\)’s head positions and work-tape contents, empty output, and state \((q,b)\), halted, or \(\ell \) according as \(C\) is in state \(q\), halted with last output bit \(1\), or halted with last output bit \(0\) (the bit \(b\) being the last output bit, \(0\) if none). Then for every such \(C\), one step of \(M^{\mathrm{acc}}\) from \(\Phi (C)\) yields \(\Phi (C')\), where \(C'\) is the result of one step of \(M\) from \(C\) (a halted configuration stepping to itself).
Let \(Q\) be a set of control states. The recognizer’s control states are the pairs \((q, b) \in Q \times \{ 0,1\} \) together with a distinguished loop state \(\ell \), and a further symbol \(\bot \) marks halting. Given \(s \in Q \cup \{ \bot \} \) (a state, or \(\bot \) for halting) and a bit \(b\), the encoded state \(\mathrm{enc}(s, b)\) is \((s, b)\) when \(s \in Q\), is \(\bot \) when \(s = \bot \) and \(b = 1\), and is the loop state \(\ell \) when \(s = \bot \) and \(b = 0\).
Let \(N\) be a nondeterministic Turing machine over the binary alphabet, let \(x\) be an input string, and let \(t \le t'\) be step budgets such that every choice word of length \(t\) drives \(N\) on \(x\) into a halted configuration. Then \(N\) accepts \(x\) within \(t'\) steps if and only if \(N\) accepts \(x\) within \(t\) steps, where acceptance within a budget means that some choice word of exactly that length yields a halted run with output the one-bit string \(1\).
Let \(M\) be a finite-state binary-alphabet Turing machine with \(k\) work tapes, state set \(Q\) and start state \(q_0\). Its halting recognizer \(M^{\mathrm{acc}}\) has the same \(k\) work tapes, states \((Q \times \{ 0,1\} ) \cup \{ \ell \} \) and start state \((q_0, 0)\). In the loop state \(\ell \) it moves no head, writes and emits nothing, and stays in \(\ell \); in state \((q, b)\) it performs \(M\)’s transition from \(q\) on the scanned symbols, except that it emits nothing, records in its bit the symbol \(M\) would emit (keeping \(b\) if none), and where \(M\) would halt it halts if the recorded bit is \(1\) and enters \(\ell \) if it is \(0\).
Let \(M\) be a finite-state binary-alphabet Turing machine and \(p : \{ 0,1\} ^{*} \to \{ 0,1\} \) a predicate such that on every input \(x\) the machine \(M\) eventually halts with the single bit \(p(x)\) as its output. Let \(M^{\mathrm{acc}}\) be the halting recognizer of \(M\), which simulates \(M\) with output suppressed while recording the last bit \(M\) emits (initially \(0\)), halts when \(M\) halts with recorded bit \(1\), and otherwise enters a loop state in which it does nothing forever. Then for every input \(x\), \(M^{\mathrm{acc}}\) halts on \(x\) (for some output \(w\) and some \(t\), it halts within \(t\) steps with output \(w\)) if and only if \(p(x) = 1\).
Let \(M\) be a finite-state binary-alphabet Turing machine and \(M^{\mathrm{acc}}\) its halting recognizer, a machine with a distinguished loop state \(\ell \) in which it moves no head, writes and emits nothing, and remains in \(\ell \). For every input \(x\), every configuration \(C\) of \(M^{\mathrm{acc}}\) on \(x\) whose state is \(\ell \), and every \(t \in \mathbb {N}\), running \(M^{\mathrm{acc}}\) for \(t\) steps from \(C\) yields \(C\) itself; in particular it never halts.
Let \(M\) be a finite-state binary-alphabet Turing machine with state set \(Q\), and let \(M^{\mathrm{acc}}\) be its halting recognizer: same work tapes, states \((q,b) \in Q \times \{ 0,1\} \) plus a stationary loop state \(\ell \), start state \((q_0, 0)\), simulating \(M\) with output suppressed and the last emitted bit recorded, halting where \(M\) halts with recorded bit \(1\) and entering \(\ell \) where it halts with bit \(0\). For every input \(x\) and every \(t \in \mathbb {N}\), the configuration of \(M^{\mathrm{acc}}\) after \(t\) steps on \(x\) has the same head positions and work-tape contents as the configuration \(C_t\) of \(M\) after \(t\) steps on \(x\), an empty output tape, and state \((q,b)\), halted, or \(\ell \) according as \(C_t\) is in state \(q\), halted with last output bit \(1\), or halted with last output bit \(0\), where \(b\) is the last output bit of \(C_t\) (\(0\) if none).
Let \(S, V\) be deterministic machines, \(d\) a halted scheduler configuration on an input \(x\), \(u\) a guessed word, and \(w\) any choice word with \(|w| = |x\, u| + 2\). Running the host under \(w\) from the guessing-phase configuration for \((d,u)\) installs the verifier-phase configuration for the initial configuration of \(V\) on the assembled word \(x\, u\).
Let \(L\) be a language over the binary alphabet such that \(x \in L\) exactly when some binary string \(u\) with \(|u| = C\, (|x|+1)^c\) has \(x\, u \in V\), and let \(M\) be a deterministic machine deciding the verifier language \(V\) within time \(A(m+1)^d\). Then there are constants \(K, r\) and a nondeterministic machine deciding \(L\) within time \(K\, (n + C(n+1)^c + 1)^r\).
Let \(S, V\) be deterministic machines and \(x\) an input. For any choice bit, one host step from the copy-state loader configuration at the right input blank with assembly buffer \(x\) yields the rewind-state configuration with the input head moved left to position \(|x|\).
Let \(S, V\) be deterministic machines, \(x\) an input, and \(w\) any word of choice bits with \(i + |w| \le |x|\). Running the reverse host under \(w\) from the copy-state loader configuration holding the first \(i\) bits of \(x\) yields the copy-state configuration holding the first \(i + |w|\) bits, one bit copied per step regardless of the choices.
Let \(S, V\) be deterministic machines and \(x\) an input with position index \(i \lt |x|\). For every choice bit, one step of the reverse host from the copy-state loader configuration holding the first \(i\) bits of \(x\) appends the bit of \(x\) at index \(i\) to the assembly buffer and advances the input head by one.
Let \(S, V\) be deterministic machines, \(d\) a halted scheduler configuration on an input \(x\), \(u\) a guessed word, and suppose \(V\) computes the single-bit output \(b\) on \(x\, u\) within \(T\) steps. Then for every choice word \(w\) with \(|w| = |x\, u| + 2 + (T+1)\), running the host under \(w\) from the guessing-phase configuration for \((d,u)\) halts with output exactly the one-bit string \(b\).
Let \(S, V\) be deterministic machines and suppose the input-normalized scheduler on an input \(x\) computes an output \(v\) within \(T\) steps while remaining live at all earlier times. Then every choice word of length \(T\) drives the reverse host from the start-of-guessing configuration to the completed-scheduler guessing configuration with some guessed word of length \(|v|\), and every word of length \(|v|\) arises this way from some choice word of length \(T\).
Let \(S, V\) be deterministic machines, \(d\) a halted scheduler configuration on an input \(x\), and \(u\) a guessed word. For any choice bit, one host step from the guessing-phase configuration for \((d,u)\) yields the assembly-rewind configuration holding the assembled word \(x\, u\) with assembly head at its last cell.
Let \(S, V\) be deterministic machines, \(d\) a scheduler configuration on an input \(x\), \(u\) a guessed word, and \(w\) a word of choice bits such that the input-normalized scheduler is still live at every time before \(|w|\). Running the reverse host under \(w\) from the guessing-phase configuration for \((d,u)\) gives the guessing-phase configuration for the \(|w|\)-step scheduler run, with the selection of \(w\) along the scheduler emission mask appended to the guessed word.
Let \(S, V\) be deterministic machines, \(d\) a non-halted scheduler configuration on an input \(x\), \(u\) a guessed word and \(b\) a choice bit. One step of the reverse host under \(b\) maps the guessing-phase configuration for \((d,u)\) to the guessing-phase configuration for the stepped configuration of the input-normalized variant of \(S\), with \(b\) appended to the guessed word exactly when that machine emits at \(d\).
Let \(S\) and \(M\) be deterministic machines, \(V\) a language, and suppose the input-normalized scheduler on an input \(x\) computes a word \(v\) within \(T\) steps while live earlier, and that \(M\) decides \(V\) within time \(A(m+1)^d\) on inputs of length \(m\). Set \(H = 2|x| + 2 + T + (|x| + |v| + 2 + (A\, (|x|+|v|+1)^d + 1))\). Then every choice word of length \(H\) makes the reverse host built from \(S\) and \(M\) halt with output the indicator bit of \(x\, u \in V\) for some guessed word \(u\) of length \(|v|\); and conversely, for every \(u\) of length \(|v|\) some choice word of length \(H\) produces a halted run with output the indicator bit of \(x\, u \in V\).
For all deterministic machines \(S, V\) and every input \(x\), the blank-tape initial configuration of the reverse host built from \(S\) and \(V\) equals the copy-state loader configuration with empty assembly buffer, input head at position one and assembly head at zero.
Let \(S, V\) be deterministic machines, \(d\) a scheduler configuration and \(y\) an assembled word. For any choice bit, one host step from the assembly-rewind configuration with head at cell \(-1\) yields the verifier-phase configuration for the initial configuration of \(V\) on \(y\), with interior boundary tag, blank verifier tapes and empty captured output.
Let \(S, V\) be deterministic machines, \(d\) a scheduler configuration, \(y\) an assembled word, \(j \le |y|\), and \(w\) any choice word with \(|w| = j + 1\). Running the host under \(w\) from the assembly-rewind configuration with head at cell \(j - 1\) reaches the verifier-phase configuration for the initial configuration of \(V\) on \(y\).
Let \(S, V\) be deterministic machines, \(d\) a scheduler configuration, \(y\) an assembled word and \(j \lt |y|\). For any choice bit, one host step from the assembly-rewind configuration with head at cell \(j\) moves the head to cell \(j - 1\), preserving \(y\) and the scheduler bank.
Let \(S, V\) be deterministic machines and \(x\) an input. From the rewind-state loader configuration with input head at the left blank and assembly buffer exactly \(x\), one host step (for any choice bit) yields the guessing-phase configuration for the initial configuration of the input-normalized scheduler on \(x\) with empty guessed word.
Let \(S, V\) be deterministic machines, \(x\) an input, \(j \le |x|\), and \(w\) any word of choice bits with \(|w| = j + 1\). Running the reverse host under \(w\) from the rewind-state loader configuration with input head at position \(j\) and assembly buffer \(x\) reaches the guessing-phase configuration for the initial normalized-scheduler configuration on \(x\) with empty guessed word.
Let \(S, V\) be deterministic machines and \(x\) an input with \(j \lt |x|\). For any choice bit, one host step from the rewind-state loader configuration with input head at position \(j+1\) and assembly buffer \(x\) moves the input head to position \(j\), leaving all tapes unchanged.
Let \(S, V\) be deterministic machines, \(x\) an input, and \(w\) any word of choice bits with \(|w| = 2|x| + 2\). Running the reverse host under \(w\) from its blank-tape initial configuration on \(x\) reaches the guessing-phase configuration for the initial configuration of the input-normalized scheduler on \(x\) with empty guessed word; all startup choices are ignored.
Let \(S\) and \(V\) be deterministic machines and let \(q\) be any control state of the reverse host built from \(S\) and \(V\) that is not a live guessing state. For every input symbol and every tuple of work-tape symbols, the two transition tables of the host agree at \(q\).
For every tape action \(a\) and every configuration \(d\) on an input \(x\), reinterpreting the result of applying \(a\) to \(d\) over the all-ones word equals applying \(a\) to the reinterpretation of \(d\); input-head clamping depends only on the input length.
Let \(M\) be a deterministic machine over the binary alphabet and suppose \(M\) computes an output \(v\) within \(t\) steps on the all-ones input of length \(|x|\). Then the input-normalized variant of \(M\) computes \(v\) within \(t\) steps on the input \(x\) itself.
Let \(M\) be a deterministic machine over the binary alphabet such that for every \(n\), on the all-ones input of length \(n\), \(M\) computes a word \(f(n)\) within \(T(n)\) steps. Then there is a function \(\tau \) with \(\tau (n) \le T(n)\) for all \(n\) such that on every input \(x\), the input-normalized variant of \(M\) computes \(f(|x|)\) within \(\tau (|x|)\) steps and is still live at every time before \(\tau (|x|)\).
For every deterministic machine \(M\) over the binary alphabet and every input \(x\), the reinterpretation of the initial configuration of the input-normalized variant of \(M\) on \(x\) equals the initial configuration of \(M\) on the all-ones word of length \(|x|\).
Let \(M\) be a deterministic machine over the binary alphabet, \(d\) a configuration and \(t\) a step count. The emission mask of the input-normalized variant of \(M\) over \(t\) steps from \(d\) equals the emission mask of \(M\) over \(t\) steps from the unary reinterpretation of \(d\).
Let \(M\) be a deterministic machine over the binary alphabet, \(d\) a configuration on an input \(x\), and \(t\) a step count. The unary reinterpretation of the \(t\)-step run of the input-normalized variant of \(M\) from \(d\) equals the \(t\)-step run of \(M\) from the reinterpretation of \(d\).
Let \(M\) be a deterministic machine over the binary alphabet and \(d\) a configuration on an input \(x\). The unary reinterpretation of one step of the input-normalized variant of \(M\) at \(d\) equals one native step of \(M\) at the unary reinterpretation of \(d\), including the absorbing halted case.
Let \(S, V\) be deterministic machines and \(d\) a halted verifier configuration. For any choice bit, one host step from the verifier-phase configuration for \(d\) halts the host and emits the single bit recording whether the complete output of \(d\) is exactly the one-bit string \(1\).
Let \(S, V\) be deterministic machines, \(d\) a verifier configuration with a valid boundary tag, and \(w\) a choice word such that the verifier is live at every time before \(|w|\). Running the host under \(w\) from the verifier-phase configuration for \(d\) yields the verifier-phase configuration for the \(|w|\)-step verifier run, with a valid tag; the choices are ignored.
Let \(S, V\) be deterministic machines, \(d\) a non-halted verifier configuration on an assembled word \(y\) carrying a valid virtual boundary tag, and fix frozen scheduler data and a physical position. For every choice bit, one host step maps the verifier-phase configuration for \(d\) to the verifier-phase configuration for the stepped verifier, with some valid updated tag; reads come from the assembly buffer and every emission of \(V\) is captured, including one on a halting transition.
Let \(S, V\) be deterministic machines and suppose \(V\) computes the single-bit output \(b\) on a word \(y\) within \(T\) steps. For every choice word \(w\) with \(|w| = T + 1\), running the host under \(w\) from the verifier-phase start on \(y\) halts with output exactly the one-bit string \(b\).
Given a scheduler configuration \(d\) on an input \(x\) and a guessed word \(u\), the corresponding configuration is the host configuration that embeds the guessing-phase invariant for \(d\) with captured word \(x\, u\) in the scheduler bank and keeps the entire verifier bank blank.
For deterministic machines \(S\) (a scheduler) and \(V\) (a verifier) over the binary alphabet, a nondeterministic machine is defined that copies its input to an assembly buffer and rewinds; runs the input-normalized scheduler, appending one nondeterministic choice bit to the assembly buffer at each scheduler emission; rewinds the assembly buffer; then runs \(V\) with its input relocated to the assembled word under a guarded virtual head, capturing every emission of \(V\); and finally halts with the single bit recording whether the captured output is exactly the one-bit string \(1\). Outside the live guessing states its two transition tables coincide.
Given a loader state index, a physical input position, a copied prefix and an assembly head value, the corresponding host configuration has every scheduler and verifier tape blank, the copied prefix on the assembly buffer, and empty physical output; the physical input is retained verbatim.
Given a completed scheduler configuration \(d\) on an input \(x\), an assembled word \(y\) and an assembly head value, the corresponding host configuration is in the assembly-rewind control state, retains the scheduler bank of \(d\) and the word \(y\) on the assembly buffer, and keeps the verifier bank blank.
Given deterministic machines \(S\) and \(V\), values for the work tapes of \(S\), an assembly value, values for the work tapes of \(V\) and a capture value, the combined assignment is the function that places the scheduler bank first with the assembly buffer as its final slot, followed by the verifier bank with the capture tape as its final slot.
For a deterministic machine \(M\) over the binary alphabet, a machine with the same states and tapes is defined whose transition table first replaces every non-blank input symbol it reads by the bit \(1\) and otherwise acts exactly as \(M\); blanks are preserved, so only interior input reads are normalized.
Given a verifier configuration \(d\) on an assembled word \(y\), a boundary tag, a physical input position and frozen scheduler tapes and heads, the corresponding configuration is the host configuration that stores the verifier state, tag and output summary in control, \(y\) on the assembly buffer serving as virtual input, the verifier work tapes in their own bank, the complete emitted output on the capture tape, and empty physical output.
Given an incoming borrow flag and a bit word, the borrow pass processes every bit of the word: a bit \(b\) met with borrow \(c\) is rewritten to \(b \oplus c\) and the borrow becomes \(c \wedge \lnot b\). It returns the final borrow flag together with the rewritten word, always retaining the full width of the word even when leading high bits become zero.
Given a list of bits, its little-endian numeric value is defined recursively: the empty word has value \(0\), and a word with first bit \(b\) and remaining word of value \(v\) has value \(2v + b\). Zero-padded words are permitted and high zero bits do not change the value.
For a binary string \(w\), the output summary is: nothing when \(w\) is empty, the bit \(b\) when \(w\) is the single-bit string \(b\), and the bit \(0\) when \(w\) has two or more bits. It is a three-valued digest that recognizes exactly the singleton outputs.
Given a current output summary \(s\) (an optional bit) and an optional emitted bit \(e\), the updated summary is \(s\) when nothing is emitted, the emitted bit when \(s\) is still empty, and the bit \(0\) when a bit is emitted while \(s\) already holds one.
For \(C, c, m \in \mathbb {N}\), the partial value \(\mathrm{split}_{C,c}(m)\) is the least \(n \in \{ 0, 1, \dots , m\} \) satisfying \(n + (C+1)(n+1)^c = m\), found by searching \(0, 1, \dots , m\) in increasing order; it is undefined when no such \(n\) exists.
For natural numbers \(C\), \(c\) and \(m\), the bounded search for a split position (which scans all \(n \le m\) for a solution of \(n + C\, (n+1)^c = m\)) returns no result if and only if no natural number \(n\) satisfies \(n + C\, (n+1)^c = m\).
For \(C, c, m \in \mathbb {N}\) let \(\mathrm{split}_{C,c}(m)\) be the least \(n \le m\) with \(n + (C+1)(n+1)^c = m\), and undefined if there is none. Then for all \(C, c, m, n \in \mathbb {N}\), \(\mathrm{split}_{C,c}(m) = n\) if and only if \(n + (C+1)(n+1)^c = m\).
Let \(N\) be a nondeterministic Turing machine over the binary alphabet that decides a language \(L\) within time \(a\, (n^c+1)\) on inputs of length \(n\), meaning that every branch halts within the budget and acceptance characterizes membership. Then for every binary string \(x\): \(x \in L\) if and only if there is a binary string \(u\) with \(|u| = 2a\, (|x|+1)^c\) such that the concatenation \(x\, u\) lies in the choice-word verifier language of \(N\) with parameters \(2a\) and \(c\).
For a nondeterministic Turing machine \(N\) over the binary alphabet with \(k\) work tapes, a deterministic machine on \(k+3\) tapes is defined that simulates \(N\) from prepared private tapes: each bit read from the choice buffer drives one step of \(N\) (with the input read from a copied virtual-input tape and every emission written to a capture tape), and the first blank on the choice buffer triggers a final transition that halts and emits one verdict bit, which is \(1\) exactly when the simulated machine has halted with accepting output summary.
Let \(N\) be a nondeterministic Turing machine over the binary alphabet and consider the embedding of a configuration of \(N\) whose choice head stands at the blank after the complete choice word \(u\). One further step of the simulation core halts with output a single bit, and that bit is \(1\) exactly when the embedded source configuration has halted with output the one-bit string \(1\).
For every nondeterministic Turing machine \(N\) over the binary alphabet and every input string \(x\), the initial configuration of \(N\) on \(x\) satisfies the virtual boundary-tag invariant with the interior tag value (its input head starts at position one, inside the window).
Let \(N\) be a nondeterministic Turing machine over the binary alphabet, let a configuration of \(N\) on input \(x\) carry a valid virtual boundary tag, and let \(u\) be a choice word with \(t \le |u|\). Then \(t\) physical steps of the simulation core of \(N\) carry the embedded configuration to the embedding of the run of \(N\) under the first \(t\) bits of \(u\), again with a valid boundary tag.
Let \(N\) be a nondeterministic Turing machine over the binary alphabet, let a configuration of \(N\) on input \(x\) carry a valid virtual boundary tag, and let \(u\) be a choice word with next unread index \(j \lt |u|\). Then one physical step of the simulation core of \(N\) transforms the embedded configuration into the embedding of the configuration obtained by one step of \(N\) under the choice bit at position \(j\), with choice index \(j+1\) and some valid updated boundary tag.
Let \(N\) be a nondeterministic Turing machine over the binary alphabet, let a configuration of \(N\) carry a valid virtual boundary tag, and let \(u\) be a choice word. From the prepared embedded configuration, the simulation core of \(N\) halts after exactly \(|u| + 1\) steps, and its output is the single bit recording whether the run of \(N\) under \(u\) from that configuration has halted with output the one-bit string \(1\).
Given a configuration of a nondeterministic machine \(N\) on input \(x\), a boundary tag, a choice word \(u\), a choice-head index \(j\), and a physical input position, the corresponding embedding is the configuration of the simulation core of \(N\) that stores the source state, tag and output summary in finite control, each source work tape on its own tape, and the words \(x\), \(u\) and the source output on the three private buffer tapes, with empty physical output.
Given an assignment of values to \(k\) source tapes and three further values, the combined assignment on \(k+3\) tapes is the function that places the source values on the first \(k\) tapes and the three extra values (a virtual input, a choice buffer, and a captured output) on the last three tapes, in that order.
Given a nondeterministic Turing machine \(N\) over the binary alphabet and natural numbers \(C\) and \(c\), the choice-word verifier language of \(N\) consists of all binary strings of the form \(x\, u\) (concatenation) where \(|u| = C\, (|x|+1)^c\) and the run of \(N\) on input \(x\) under choice word \(u\) has halted with output exactly the one-bit string \(1\).
Let \(N\) be a nondeterministic Turing machine over the binary alphabet, let \(C, c\) be natural numbers, and let \(x\) and \(u\) be binary strings with \(|u| = C\, (|x|+1)^c\). Then the concatenation \(x\, u\) belongs to the choice-word verifier language of \(N\) with parameters \(C, c\) if and only if the run of \(N\) on input \(x\) under choice word \(u\) has halted with output exactly the one-bit string \(1\).
Let \(N\) be a nondeterministic Turing machine over the binary alphabet and let \(C, c\) be natural numbers. If no natural number \(n\) satisfies \(n + C\, (n+1)^c = |y|\) for a binary string \(y\) (so the bounded split search fails at length \(|y|\)), then \(y\) does not belong to the choice-word verifier language of \(N\) with parameters \(C, c\).
Let a machine run from a live configuration to a halted configuration in \(t\) steps. Then there is a strictly positive time \(u \le t\) at which the run first halts, and the configuration at \(u\) equals the configuration at \(t\); padding by absorbed halted steps changes no endpoint field.
If a single-bit test \(p\) and two string functions \(f\) and \(g\) are polynomial-time computable, then so is the function returning \(f\) of the input when the test holds and \(g\) of the input otherwise; the three budgets fit one common maximal degree.
For all parameters \(C\) and \(e\), every instance \(x\), and every assignment agreeing with \(x\) on the first \(|x|\) variables, there is a certificate \(u\) of length exactly \(C(|x|+1)^e\) such that the recovered input word of the first \(|x| + C(|x|+1)^e\) variables is the concatenation of \(x\) and \(u\); the certificate polynomial is not enlarged and no bounded-length witness is substituted.
A machine satisfies the clean contract for a result map, an emission map, and a budget when it has a positive tape count and, for every native input and argument word, it first halts at some positive time within the budget of the argument length, at the canonical seam holding the result word and having emitted exactly the prescribed chunk.
Every polynomial-time computable string function \(f\) has a first-halting clean module that preserves the argument word, emits exactly \(f\) of the argument as output, and runs within a polynomial budget of the argument length, including cleanup and the final halt.
Every polynomial-time computable string function \(f\) has a first-halting clean module that replaces the argument word by \(f\) of the argument, emits nothing, and runs within \(A(n+1)^d\) steps of the argument length for some constants \(A, d\).
If a machine has, for every input and argument, a positive first exit within a budget \(B\) at the canonical seam with a result word and emitted chunk, then its release wrapper satisfies the uniform clean contract with budget \(B + 1\); the minimal-halt cut retains the entire clean configuration.
For every four words, the verdict-replaying comparator computes, from the packed encoding of the four words, the single bit deciding whether the value of the first word plus the second equals the value of the third plus the fourth, within \(35\) times the encoded argument length plus one.
The extended comparator machine behaves as the natively loaded four-word comparator, except that on reaching a completed verdict control it emits that single verdict bit and physically halts; the source comparator and its signed arithmetic are unchanged.
Let an oblivious machine decide, within time \(T\) and for every word \(y\) of length \(|x| + C(|x|+1)^e\), the indicator of membership of \(y\) in a verifier language \(V\). If some certificate \(u\) of length exactly \(C(|x|+1)^e\) makes the concatenation of \(x\) and \(u\) belong to \(V\), then the Cook–Levin tableau of \(x\) at horizon \(T\) is satisfiable: the literal input assignment together with the raw encodings of the genuine snapshots satisfies every family, and the singleton true output supplies the no-false clauses.
For every input word \(x\), the release-wrapped native input copier, started over the blank state word, first halts at a positive time at most \(2|x| + 3\) at the canonical seam whose stored word is exactly \(x\), with silent output and unchanged heads.
For a counter coordinate \(i\), a persistent body word, and a time \(t\), the stored count is the bounded value, under the stored horizon, of the trajectory field at index \(t\) times the field count plus \(i\), read from the actual chronological trajectory.
For every source machine, parameters, instance \(x\), cursor \(j\), time \(t\) at most the producer horizon, and coordinate \(i\), the stored count read from the body word equals the canonical movement count at time \(t\) over the all-false reference input; the bound never truncates valid data.
For every counter coordinate and polynomial-time unary time producer, the unary word of the stored movement count at the produced time is polynomial-time computable; projections, clock construction, repeated field scans, and decoding calls are all charged.
There are a clean module and constants such that the module replaces its argument word by the pair encoding of the empty word with that argument, installing an empty unary emission cursor in front of the packed data, while emitting nothing; neither the header nor either stored table changes.
For every source machine, parameters, instance \(x\), and cursor \(i\), the stored input length is \(m = |x| + C(|x|+1)^e\), the stored horizon is the producer horizon \(T\), the stored trajectory is the inclusive serialized trajectory of \(T+1\) rows over the all-false input of length \(m\), and the stored visit table is the ordered table of \(T+1\) visit rows.
Let a machine compute, on input \(y\) within time \(T\), the single indicator bit of membership of \(y\) in a language \(V\). Then no snapshot before time \(T\) emits the bit false exactly when \(y \in V\); silent or junk traces cannot satisfy this equivalence, because the total decider supplies its actual verdict bit.
If a binary-word producer and a unary bound producer are polynomial-time computable, then so is the function returning the unary value of the produced binary word when that value is at most the produced bound, and the empty word otherwise; all malformed requests retain the proved polynomial cost.
For a binary target word \(w\) and a candidate count \(j\), the decoding state retains \(w\), the unary candidate of \(j\) tokens, and the most recent successful unary answer: the unary value of \(w\) when it is at most \(j\), and the empty word otherwise.
One decoding step yields the state whose unary candidate is extended by one token; the genuine binary encoding of the new candidate’s length is compared with the retained target, and the candidate is recorded as the answer exactly on equality, the old answer being kept otherwise.
If a unary value is polynomial-time computable, then for every fixed constant \(k\) the unary word of its quotient by \(k\) is polynomial-time computable: a finite-control counter emits one bit per completed block, and division by zero yields the empty word.
If a word producer and a length-witness producer are polynomial-time computable, then so is the function dropping, from the produced word, a number of leading bits equal to the witness’s length; each tail call is charged even after the word is exhausted.
On a persistent body word, the emitter outputs the current fragment while the cursor is at most the stored last round, appending the single formula-terminator bit exactly at the last round, and outputs nothing in the saturated state; empty group fragments still have their own positive round.
For every source machine, state code, parameters, instance \(x\), and cursor \(j\), the chunk selector on the body word at \(j\) equals the chunk of round \(j\) of the pure tableau group list when \(j\) is at most the last round, and the empty word otherwise; member counts and the sole-last-terminator convention are preserved.
If two word producers are polynomial-time computable, then the single-bit test deciding whether their outputs denote equal natural numbers is polynomial-time computable, using the banked four-word comparator with the two remaining words empty.
Let an oblivious machine decide a verifier language \(V\) within time \(c(A(n+1)^d+1)^2\). For every instance \(x\), the Cook–Levin tableau of \(x\) at the producer horizon is satisfiable exactly when some certificate \(u\) of length \(C(|x|+1)^e\) makes the concatenation of \(x\) and \(u\) belong to \(V\).
On a persistent body word, the selector emits, by successive guarded offsets of the current cursor: a pinning unit for cursors inside the retained instance; the initial template at the post-pin ordinal zero; a state template; a stored-coordinate input fragment; a stored-visit work fragment; an acceptance template; and the empty word past all families. Only the current group’s fragment is emitted; the formula terminator belongs to the separate last-round check.
For every source machine, state code, parameters, instance \(x\), and cursor \(j\), the native fragment selector on the body word at \(j\) agrees word for word with the serialized fragment of the \(j\)-th member of the unchanged pure tableau group list at the producer horizon; the guarded offsets are precisely list lookup, and a work ordinal’s quotient and remainder name its time and tape in the original order.
For every source machine and state code, the complete ordered fragment selector is one total polynomial-time transduction: every numeric operation is unary bounded, stored binary counters are decoded only within the explicit unary horizon, and templates and tape choices are fixed finite data of the machine.
For every source machine and parameters, the last-round function of the input length is polynomially bounded, and the function sending an input to the binary representation of its last-round value is polynomial-time computable; the fuel is the last round index, read from an actual unary producer followed by the native binary length machine.
For every source machine, state code, template kind, wiring data, and assignment \(a\), the wired clause group evaluates to true under \(a\) exactly when the bit-level meaning of the template holds: the wiring exposes the raw source and target blocks and the selected input bit exactly, so satisfaction means bitwise equations, never merely equality of decoded snapshots.
For a fixed machine, state code, and template kind, if the unary input count, source time, target time, and selected index are polynomial-time computable, then the serialized fragment of the wired clause group is polynomial-time computable; only the unary wiring varies with the input.
Given parameters \(n\), \(k\), \(T\) and the six clause-family generators, the selector at ordinal \(i\) subtracts the family offsets in order: the \(n\) pin groups, the single initial group, \(T\) state groups, \(T+1\) input groups, \(k(T+1)\) work groups in time-then-tape order via quotient and remainder, and \(T\) accept groups, with the empty formula past the end.
Suppose a module runs from its start over an argument word to a halted canonical seam with a result word and output chunk in time \(t\), never halting earlier. Then the host, started at that module’s entry over the shared bank with an output prefix, never shows the loop anchor before \(t\), and at \(t\) sits at the module’s resolved return over the result word with the output prefix extended by exactly the module’s chunk.
If the module at index \(i\) of the host satisfies the uniform clean contract for result map \(f\), emission map \(g\), and budget \(B\), then from the host entry of \(i\) over any argument and output prefix, the host reaches, within \(B\) of the argument length and without touching the loop anchor, the resolved return over \(f\) of the argument with the output extended by exactly \(g\) of the argument.
For every module with a positive tape count, control, stored word, output prefix, and optional state, forwarding the padded clean configuration yields the host’s canonical seam over the shared bank’s state word, with output exactly the prefix followed by the module’s output.
Over one shared padded work bank, the host runs the five modules through their call chain: the anchor dispatches into the emission module, and each module’s transitions are forwarded with halts redirected to the module’s resolved return; the clean-call contracts restore all scratch before each next call.
Let \(y\) be a word of length \(m\) and let the assignment agree with \(y\) on the first \(m\) variables. Then the initial meaning holds exactly when the raw block at time zero equals the encoding of the machine’s genuine initial snapshot on \(y\), with the empty-input case included.
Let \(y\) be a word of length \(m\) and let the assignment \(a\) agree with \(y\) on the first \(m\) variables. Then the input clause group at time \(t\) evaluates to true under \(a\) exactly when the input slice of the raw block at \(t\) encodes the input symbol of \(y\) at the oblivious input position of time \(t\).
For a persistent body word and a time \(t\), the input fragment serializes the present-input table wired to input variable \(p - 1\) when the stored input position \(p\) is interior, with \(0 \lt p\) at most the stored input length, and the blank-input table with a dummy index otherwise.
For every source machine, state code, parameters, instance \(x\), cursor \(j\), and time \(t\) at most the producer horizon, the stored-coordinate input fragment equals the serialized fragment of the pure input clause group at time \(t\) over the reference length.
For every source machine, state code, and polynomial-time unary time producer, the serialized input fragment at the produced time is polynomial-time computable; both template choices are finite protected tables, and the branch predicate and shifted variable address are charged.
For every source machine, parameters, instance \(x\), cursor \(j\), and time \(t\) at most the producer horizon, the stored shifted input position equals the public oblivious input position at time \(t\) over the reference length \(|x| + C(|x|+1)^e\).
Let \(f\), an initial-word producer \(s\), and a unary count producer \(N\) be polynomial-time computable, and suppose every iterate of \(f\) up to \(N\) of the input has length at most \(K((|s| + N + 1))^2\) in the produced values. Then the function sending an input \(x\) to the \(N(x)\)-fold iterate of \(f\) at \(s(x)\) is polynomial-time computable; each iteration is one clean call inside the real unary repetition host.
If a polynomial-time computable step function never increases word length, then for any polynomial-time initial-word and unary-count producers, the counted iterate is polynomial-time computable; repeated sequential projections and tail scans therefore need no random access.
For every finite transducer, input split into a prefix and an unread suffix, state, and emitted output, running the machine for the suffix length plus one steps halts, and its output is exactly the prior emitted output followed by the recursive transduction of the suffix from the current state.
For a finite state set with step, emission, and finishing functions and a start state, the transducer is a machine with no work tapes that scans its input left to right, emits the per-bit outputs, and halts at the right boundary after the optional final emission.
For a finite state set with a step function, an optional per-bit emission, and an optional final emission, the transduction of a word processes each bit in order, appending its optional emitted bit and updating the state, and appends the optional final bit at the end; at most one bit is emitted per input bit.
For a template kind, wiring data \(m\), source time \(s\), target time \(t\), selected index \(p\), and an assignment \(a\), the meaning states: for the initial kind, that the raw target block equals the encoding of the start snapshot reading bit \(a(p)\) or a blank; for the state kind, that the target state slice encodes the stepped state of the decoded raw source block; for the input kind, that the target input slice encodes the selected bit or a blank; for a work kind, that the target work slice encodes the symbol written or kept by the decoded source at its last earlier visit, or a blank; and for the accept kind, that the decoded source does not emit false.
One cursor update on the persistent body word extends the unary cursor by one token while the cursor is at most the stored last-round index, and leaves it unchanged beyond that saturation point; the complete protected producer answer is untouched.
For every source machine, parameters, instance \(x\), and cursor \(i\) at most the last round plus one, the native cursor update of the body word at \(i\) is the body word at \(\min (i+1, R+1)\), where \(R\) is the last round; the saturated cursor is stationary.
For every machine, input \(y\), and horizon \(T\), the append-only output after \(T\) steps contains no false bit exactly when no snapshot before time \(T\) emits the bit false; the equivalence is an exact word statement with no halting assumption.
Let an emission function on body words be polynomial-time computable and agree, on every body word at a legal cursor \(i\), with the chunk of round \(i\) of the tableau group list, and with the empty word past the last round. Then the function sending an instance \(x\) to the serialization of the pure Cook–Levin tableau of \(x\) at the producer horizon is polynomial-time computable: startup, cursor updates, clean first returns, saturated silence, exact fuel, and one common polynomial budget are all discharged, and no satisfiability fact enters the word identity.
For every source machine, state code, and parameters, the function sending an instance \(x\) to the serialization of the pure Cook–Levin tableau of \(x\) at the producer horizon is computable by one finite machine in polynomial time from genuine native input. No certificate or satisfiability assumption is involved.
For a source machine, parameters, instance \(x\), and cursor value \(i\), the persistent body word is the pair of the unary cursor of \(i\) tokens with the entire unchanged packed producer answer; the original instance remains inside its header.
For every source machine and parameters, there are a clean module and constants such that the module replaces its argument word, the genuine retained native input, by the complete packed preparation data, header, trajectory, and visit table, emitting nothing, within a polynomial budget; the completed producer itself is consumed, not substituted.
For every machine with \(k \le K\) tapes, source configuration, and duration \(t\), running the padded machine for \(t\) steps from the padded configuration equals padding the source’s \(t\)-step run; the padded source observes exactly its original work symbols.
The padded form of a configuration on \(k\) tapes is the configuration on \(K\) tapes that keeps its state, input position, and output, places its tapes and heads on the first \(k\) slots, and leaves every additional tape genuinely blank with head zero.
Let the machine be oblivious, let \(y\) have length \(m\), and suppose an assignment’s raw blocks satisfy: the block at time zero is the encoded initial snapshot; the state meaning at every \(t \lt T\); the input-slice equations at every \(t \le T\); and the work-slice equations at every \(t \le T\) on all tapes. Then for every \(t \le T\) the raw block at \(t\) equals the encoding of the genuine snapshot of the run on \(y\) at time \(t\); junk blocks are excluded bit for bit.
For every language \(L\) in \(\mathrm{NP}\), the map sending an instance \(x\) to the serialization of its Cook–Levin tableau at the producer horizon is a polynomial-time many-one reduction from \(L\) to \(\mathrm{SAT}\); neither the certificate parameters nor any public statement is modified.
Let the emission module satisfy the clean contract preserving its argument and emitting a chunk, and the update module the clean contract replacing the argument and emitting nothing. Then from the anchor over any body word \(w\), the host first returns to the anchor at a positive time at most \(1\) plus both budgets at \(|w|\), having appended exactly the emitted chunk and replaced \(w\) by its update; empty chunks and a stationary saturated cursor are covered.
For every machine, input \(y\), and horizon \(T\), the physical output after \(T\) steps is the chronological concatenation of the optional emissions of the snapshots at times \(0, \dots , T-1\), including the write on a halting transition.
For a finite list of polynomial-time computable word functions and a polynomial-time unary selector, the function applying the selected list entry to the input, with the empty-word constant as out-of-range default, is polynomial-time computable; no unbounded transition table is consulted.
Let an oblivious machine decide, within time \(T\) and for every word \(y\) of length \(|x| + C(|x|+1)^e\), the indicator of membership of \(y\) in a verifier language \(V\). If the Cook–Levin tableau of the instance \(x\) at horizon \(T\) is satisfiable, then there is a certificate \(u\) of length exactly \(C(|x|+1)^e\) with the concatenation of \(x\) and \(u\) in \(V\). The raw-block reconstruction excludes junk assignments before any decoder use, and totality plus the no-false family forces the accepting verdict.
For every source machine and parameters, there are install and cursor machines with polynomial constants such that, for arbitrary emission and update modules, the five-module host started from its native initial configuration reaches, within the summed startup budget and without ever showing the loop anchor earlier, the anchor over the persistent body word with cursor zero: the complete packed producer answer behind an empty unary cursor. Startup is silent and never executes the emission or update modules.
Suppose a clean call runs from its entry over an argument word to its exit over a result word with some emitted output, in positive time \(t\), never showing the exit strictly between \(0\) and \(t\). Then the release wrapper, started unreleased at the entry, halts physically at time \(t + 1\) with exactly the result’s canonical seam, the same emitted output, and no further transition.
Given a machine with entry and exit controls, the wrapped machine carries one release flag: it executes the underlying transitions, raising the flag after the first step, and physically halts exactly when the flag is set and the control is the exit; an entry equal to the exit therefore still executes at least one step.
For every source machine, parameters, instance \(x\), and cursor \(i\), the retained instance of the persistent body word is \(x\) itself, and the stored last-round index equals the last round of the schedule at input length \(|x|\), tape count, and producer horizon.
From the persistent body word, the stored last-round index is computed from the retained instance’s length, the machine’s tape count, and the length of the stored unary horizon, header field five; the horizon is a unary length, not an unproved binary-to-unary conversion.
For every source machine, state code, parameters \(C\) and \(e\), instance \(x\), horizon \(T\), and assignment \(a\), with \(m = |x| + C(|x|+1)^e\): the tableau evaluates to true under \(a\) exactly when the first \(|x|\) variables are pinned to the bits of \(x\); the initial meaning holds; the state meaning holds at every \(t \lt T\); every inclusive input group with \(t \le T\) evaluates to true; every inclusive work group with \(t \le T\) evaluates to true on every tape; and the acceptance meaning holds at every \(t \lt T\).
For every fixed CNF table and every wiring whose unary index is polynomial-time computable at each of the finitely many template variables, the function emitting the serialized fragment of the relabeled table is polynomial-time computable; every literal occurrence pays its variable-index cost, and the formula terminator is deliberately absent from the fragment.
For every machine, state code, input word \(y\), and time \(t\), the raw block of the trace assignment at time \(t\) equals the encoding of the genuine snapshot of the run on \(y\) at time \(t\); quotient and remainder recover each complete block.
For a machine with a state code and an input word \(y\), the trace assignment maps the first \(|y|\) variables to the literal bits of \(y\), and each later variable to the corresponding bit, located by quotient and remainder modulo the block width, of the raw encoding of the chronological snapshot of the run on \(y\).
For every source machine, parameters, instance \(x\), cursor \(j\), time \(t\) at most the producer horizon, and tape \(\tau \), the stored visit code equals the optional-time code of the public greatest strictly earlier visit; the reference machine is never rerun and the protected table is never replaced.
For every source machine, parameters \(C\) and \(e\), instance \(x\), and count \(N\), the ordered visit table of \(N\) rows equals the generic sequential row stream whose row at time \(t\) lists, per work tape, the optional-time code of the greatest strictly earlier visit over the reference length \(|x| + C(|x|+1)^e\); time zero is included.
For every element family on times and tapes, tape count \(k\), bound \(N\), and index \(j \lt kN\), the \(j\)-th entry of the list enumerating, for each time below \(N\), the \(k\) tape entries in order, is the entry at time \(j\) divided by \(k\) and tape \(j\) modulo \(k\); the zero-tape case has no legal index.
For every input count \(m\), time \(t\), tape \(\tau \), and assignment \(a\), the work clause group evaluates to true under \(a\) exactly when the work slice of the raw block at \(t\) is the encoding of a blank on a first visit, and otherwise of the symbol written or kept, at the greatest strictly earlier visit \(s\), by the snapshot decoded from the raw block at \(s\).
For every source machine and state code, the work-family tape choice is polynomial-time computable: finite selection over the per-tape fragment functions, driven by the native unary quotient and remainder of the work ordinal, including the empty zero-tape family.
For a work tape \(\tau \), a persistent body word, and a time \(t\): an empty stored visit code selects the blank first-visit table; a nonempty code selects the earlier-visit table wired from the source time decoded, bounded by the stored horizon, from the code’s payload. A present visit at time zero has its distinct nonempty marker.
For every source machine, state code, parameters, instance \(x\), cursor \(j\), time \(t\) at most the producer horizon, and tape \(\tau \), the stored-visit work fragment equals the serialized fragment of the pure work clause group at time \(t\) and tape \(\tau \) over the reference length; the stored earlier source is strictly below the target, so bounded decoding preserves every legitimate visit, including source zero.
For every input word, input position, selection flag, counter value \(n\), and width bound \(W\) with the binary representation of \(n\) at most \(W\) bits long, running one counter component for \(2W + 2\) steps from its selection-dependent start control ends in the return control at head zero, holding the binary representation of \(n+1\) if the component was selected and of \(n\) otherwise.
For every input word, input position, selection vector, counter values, and common width bound \(W\) covering all the values’ binary representations, running the counter bank for \(2W + 2\) steps increments exactly the selected coordinates, leaves the unselected ones unchanged, and restores every head to zero; the cost is independent of the number of counters.
For every input word, input position, selection vector, counter values, and common width bound \(W\), there is a time \(t \le 2W + 2\) such that the bank control is not all-returned at any earlier time, and at \(t\) the bank holds exactly the requested incremented values with every head restored to zero. The duration may be zero when nothing is selected; positivity is supplied by the surrounding dispatch controller.
For every input position, every family of component configurations, and every coordinate \(i\), the bank’s component action at \(i\) computed from the product configuration equals the one-tape counter transition of component \(i\) on its own scanned symbol, or the stationary action when component \(i\) has stopped.
For every input position, family of component configurations, and duration \(t\), running the counter bank for \(t\) steps yields the product of the \(t\)-step runs of all components, so every counter trajectory is an exact projection of the one native bank run.
For every input position and family of component configurations, one step of the counter bank equals the product of one single-tape counter step per component: tape contents and independently positioned heads are each updated by their own component, the physical input stays fixed, and all component outputs are suppressed.
Given an input position and one single-tape counter configuration per coordinate, the bank configuration has the vector of component controls as its state, each component’s tape and head on the corresponding bank tape, the supplied physical input position, and empty output: every increment is silent.
Given a vector of optional component controls and the scanned symbols of \(l\) counter tapes, the component action at coordinate \(i\) is the increment machine’s transition for control \(i\) on symbol \(i\) when that component is live, and a stationary no-write action when the component is stopped.
For a fixed number \(l\) of disjoint binary-counter tapes, one machine performs simultaneous in-place increments: its finite control holds the vector of component controls, each step applies every component’s counter action to its own tape, and a returned component stays idle while others are still carrying or rewinding.
If two snapshot blocks of \(B + 2 + 2k\) bits have equal state slices, equal input slices, and equal work slices on every tape, then the blocks are equal bit for bit. Equality of the decoded fields alone would not suffice; the claim pins every individual bit.
Given a machine with \(k\) work tapes and a binary code of width \(B\) for its optional states, a snapshot, consisting of a state, an input symbol, and one symbol per work tape, is encoded as a vector of \(B + 2 + 2k\) bits: the state bits, then two input-symbol bits, then two bits per work tape, in literal consecutive slices.
For every word \(w\) and bit \(b\), updating the canonical buffer tape of \(w\) at cell \(|w|\), its right blank, with \(b\) yields the canonical buffer tape of \(w\) extended by \(b\); no other cell changes, including negative cells and the empty-buffer boundary.
Given a last round index \(R\), a family of clause groups indexed by round, and a round \(i\), the chunk at \(i\) is the serialized fragment of the group at \(i\), followed by the single formula-terminator bit exactly when \(i = R\); an empty final group still emits the terminator.
Let \(R\) be the last round index and let a family assign a clause group to each round. For every \(j \le R\), the concatenation of the chunks of rounds \(0,\dots ,j-1\) equals the serialized fragment of the concatenation of the first \(j\) groups, so no formula terminator appears in a proper prefix of the rounds.
Let \(R\) be the last round index and let a family assign a clause group to each round. The concatenation of the chunks of rounds \(0,\dots ,R\) equals the full serialization of the concatenation of all \(R+1\) groups, formula terminator included. The identity concerns output words only and makes no satisfiability claim.
For every input word, input position, and four words, one comparator step at the common right boundary reads the joint blank, stores the final carry verdict in finite control, and starts rewinding by moving all heads left; this includes the case where all four words are empty.
For every input word, input position, four words of length at most \(W\), there is a time \(t\) with \(0 \lt t \le 2W + 2\) and a bit \(b\), true exactly when the value of the first word plus the second equals the value of the third plus the fourth, such that no completed control appears before \(t\) and at \(t\) the comparator sits in the completed control carrying \(b\) with all four words unchanged and heads at zero.
For every input word, input position, four words, and \(j\) at most the common width, running the comparator for \(j\) steps from the start reaches the scanning control after \(j\) columns with all heads at \(j\); every prefix of the forward scan is an actual aligned native run.
For every input word, input position, four words, and position \(j\) strictly below the common width, one comparator step from the scanning control after \(j\) columns yields exactly the next column: the control is the orbit value after \(j+1\) columns and all four heads have moved right.
For every input word, input position, four words, stored verdict, and count \(j\) at most the common width, running the comparator in its rewind control at aligned position \(j-1\) for \(j+1\) steps returns all four heads to zero and enters the completed control, preserving every word and the verdict.
For four words, the comparator control starts with clear carries and a set equality flag, and after \(j+1\) columns is the register update applied to the control after \(j\) columns, reading the four bits at position \(j\) with blanks treated as absent.
For every four words and every column count \(j\), the comparator invariant at the control after \(j\) columns, applied to the words with their first \(j\) bits dropped, holds exactly when the value of the first word plus the second equals the value of the third plus the fourth; columns beyond a short word read zeros, never uncharged random access.
For comparator registers and four remaining high words, the invariant states that the equality flag is still set and that the value of the first pair of words plus the first carry equals the value of the second pair plus the second carry, given that all earlier low digits matched.
A machine with four work tapes scans four read-only words in lockstep, feeding each column through the finite addition tables, then rewinds all four heads to zero and idles in a live control carrying the Boolean verdict. No word is modified and nothing is physically emitted.
Given three Boolean registers, two running carries and an equality flag, and the four scanned column bits, the updated registers are obtained by feeding each pair of bits through its own addition column: the two outgoing carries are stored, and the equality flag is conjoined with the comparison of the two low result bits.
For every comparator registers \(q\) and four words, the comparator invariant on the full words holds exactly when it holds for the updated registers on the four tails: equal sums are equivalent to equal low bits together with equal remaining high sums, and unequal low bits have opposite parity, which no later column can repair.
Let a machine compute output \(y\) from input \(x\) within time \(a\), and a second machine compute \(z\) from \(y\) within time \(b\). Then the buffered composition of the two machines computes \(z\) from \(x\) within time \(2a + b + 2\); the second machine’s contract is needed only on the actual intermediate word.
For every input word, input position, source word \(w\), and record, there is a time \(t\) with \(0 \lt t \le 3|w| + 3\) such that the copier is never in the completed control before \(t\), and at \(t\) it holds the record extended by the self-delimiting encoding of \(w\), with the source untouched and both heads as specified; a host can therefore dispatch on completion without knowing the field length.
For every input word, input position, source word split into a prefix \(u\) and suffix \(w\), and record, running the copier for \(2|w|\) steps from the reading control at cell \(|u|\) appends the doubling of every bit of \(w\) to the record and leaves the source head at the right end; the field separator has not yet been written.
For every input word, input position, source prefix \(u\), remaining source bits starting with \(b\), and record, two steps of the copier from its reading control at cell \(|u|\) append the two bits \(b, b\) to the record and advance the source head to cell \(|u| + 1\); the source tape is unchanged.
For every input word, input position, source word \(w\), record, and count \(j \le |w|\), running the copier in rewinding control at source cell \(j - 1\) for \(j + 1\) steps returns the source head to cell zero and enters the completed control, preserving the entire accumulated record; each crossed bit costs one transition and leaving the left blank costs one more.
For every input word, input position, source word \(w\), and record, running the copier for \(3|w| + 3\) steps from the start yields the completed control, in which the record is extended by exactly the self-delimiting pair encoding of \(w\) with empty payload and the source head is restored to zero; the empty field still takes three positive transitions.
For every input word, input position, source word \(w\), and record, two steps of the copier from the reading control at the source’s right blank yield the rewinding control, with the aligned separator bits false, true appended to the record and the source head moved one cell left.
For every copier configuration and every action that writes a bit \(b\) at the record head while moving the source head by \(d\), applying the action yields the configuration with exactly \(b\) appended to the record and the source head moved by \(d\); nothing else changes.
For an input word, a control state, an input position, a source-head cell, a source word, and an accumulated record, the copier configuration holds the source word and the record on its two buffer tapes, with the record head at the record’s right end.
A five-state machine with two work tapes copies a self-delimiting field: it doubles each source bit onto the record tape, writes one aligned false-true separator, rewinds the source head to cell zero, and stops in an absorbing completed control with the record head at its new right blank. The physical output stays silent and the source tape is read-only.
For every input word, input position, prefix \(u\), and suffix \(v\), running the increment machine in carry control at cell \(|u|\) for one more step than the carry count of \(v\) replaces \(v\) by its increment and leaves the machine in rewind control one cell left of the last written position. Exactly the initial true bits of \(v\) are cleared before the final true bit is written.
For an input word, an input position, and a stored word split as a prefix \(u\) followed by a suffix \(v\), one step of the increment machine in carry control at cell \(|u|\) yields: when \(v\) starts with a true bit, the cell rewritten to false with the head moved right, still in carry control; and otherwise the cell rewritten to true with the head moved left, in rewind control. Only the scanned cell changes.
For every input word, input position, and stored word \(w\), there is a time \(t\) with \(0 \lt t \le 2|w| + 2\) such that the increment machine, started in carry control at cell \(0\), is never in the return control before \(t\) and at time \(t\) sits in the return control at cell \(0\) with the stored word replaced by its increment.
For every input word, input position, stored word \(w\), and count \(j \le |w|\), running the increment machine in rewind control at cell \(j - 1\) for \(j + 1\) steps crosses the written cells, detects the untouched blank at cell \(-1\), and ends at cell \(0\) in the absorbing return control, with the stored word unchanged.
For every input word, input position, and stored word \(w\), running the increment machine from carry control at cell \(0\) for exactly twice the carry count of \(w\) plus two steps ends in the return control at cell \(0\) with the stored word replaced by its increment; the native input head never moves from its supplied position.
For all naturals with \(R \le T \le n\) and a recording time within the standard recording ledger, the recording time plus \(R + 4\) is at most \((10 + 10k + 8l + 7)(n+1)^2\): the same argument envelope pays for returning one final count instead of the full trajectory, including its physical output and halt.
For a prescribed input-head move and a scanned work symbol, the carry action is as follows: on reading a true bit it writes false, moves the work head right, and keeps carrying; on any other symbol it writes true, moves the work head left, and switches to the rewinding state.
For an input word, a control state, an input position, a work-head cell, a stored word, and an output, the canonical configuration of the increment machine holds the stored word on its single work tape; these configurations carry the carry, rewind, and return invariants.
The increment of a little-endian bit list is as follows: a leading false bit turns into true; a leading true bit becomes false and the carry propagates into the rest; the empty list becomes the single bit true, extending the word on overflow.
For every source machine, counter coordinate, virtual input \(y\), horizon \(T\), and retained fields, the count-output machine computes the binary representation of the selected movement count at time \(T\) from the packed layout, and its time bound is the quadratic envelope in the encoded argument length that also covers the full trajectory output.
For a selected movement-counter coordinate, the count-output machine runs the prepared recorder host to completion and then replays that counter’s final canonical binary word as the physical output, starting from its restored head at zero.
For a source machine and a virtual input \(y\), the movement counts at time \(0\) are all zero, and the counts at time \(t+1\) advance the counts at time \(t\) by the actual clamped displacements of the source step at time \(t\), computed with the deterministic tag. Counts record actual moves, not requested outward moves, and freeze after the source halts.
For every source machine, virtual input \(y\), time \(t\), and coordinate, the signed value of the stored count pair equals the reference position of the configuration reached at time \(t\): the shifted clamped input position on the first coordinate and the signed work-head positions on the rest.
For every oblivious machine, reference length \(m\), and time \(t\), the signed value of the stored input-head count pair over the all-false reference input, plus one, equals the public input position at time \(t\), and the signed work count pairs equal the public work positions, including the shifted input origin.
For every source machine, virtual input \(y\), time \(t\), and valid tag \(b\) for the position reached at time \(t\), advancing the counts at time \(t\) by the displacements computed from \(b\) yields exactly the counts at time \(t+1\).
For bit lists \(u\) and \(v\) and a bit \(b\), updating the tape that holds \(u\) followed by \(v\) at cell \(|u|\) with the symbol \(b\) yields the tape holding \(u\) followed by \(b\) and then the tail of \(v\): the write either replaces the first suffix bit or extends the word, all other cells are unchanged, and negative cells stay blank.
A four-state machine with one work tape increments a little-endian binary counter in place: one state propagates the carry, one rewinds the head to the origin, and an absorbing live state marks completion. The machine is silent: it never advances its native input, never emits output, and never starts a second increment after returning.
Let \(R\) be a bound, let a packing map assign a word to each cursor value, and suppose a step function sends the packed word of \(i\) to the packed word of \(\min (i+1, R+1)\) for every \(i \le R+1\). Then for all \(i\), iterating the step function \(i\) times from the packed word of \(0\) yields the packed word of \(\min (i, R+1)\): the cursor reaches each legal round and then stalls at the one-past-last value.
Suppose a native machine with a distinguished anchor state, a packing of cursor values into words, abstract step and emission functions, a round bound \(R\), and a budget \(P\) are given, where \(R\) and \(P\) are polynomially bounded and some helper machine computes the binary representation of \(R\) within time \(P\). Assume: stepping a packed cursor advances it saturatingly up to one past the last round; emission at cursor \(i \le R\) produces exactly the chunk of round \(i\), and nothing beyond the last round; from any input the machine reaches the packed start word, first returning to the anchor, within time \(P\); and from each packed cursor the machine returns to the anchor in positive time at most \(P\), having appended exactly the prescribed chunk to its output. Then the function sending an input to the serialization of the concatenation of all clause groups over all rounds is computable in polynomial time. No preparation producer or native body is constructed here.
Every nonempty finite type carries a binary code whose width is the cardinality of the type: since a finite set injects into the Boolean vectors indexed by itself, an embedding exists, and any left inverse of it serves as the total decoder. Taking the cardinality as bit width avoids any logarithm convention.
A binary code for a type consists of a bit width \(B\), an encoding map from the type into \(B\)-bit vectors, and a total decoding map that is a left inverse of the encoding. The code is fixed once for the machine, before any reduction input.
For a fixed bit \(b\), an input \(x\), a position \(i \le |x|\), and a current output word \(w\), running the constant-fill machine for \(|x| - i + 1\) steps from the configuration scanning position \(i\) with output \(w\) halts at the right boundary with output \(w\) extended by \(|x| - i\) copies of \(b\).
If a reference-input producer and a unary-horizon producer are polynomial-time computable, then for every counter coordinate the function sending an input to the binary representation of that movement count at the produced horizon is polynomial-time computable; keeping the time unary avoids any exponential clock.
Let a machine run from a start configuration to a finish configuration in \(B\) steps, where the finish control is a distinguished absorbing return state, every configuration in that state is fixed by one step, and the start is not in the return state. Then there is a strictly positive time \(t \le B\) at which the run first reaches the return state, and the configuration at \(t\) is already the exact finish configuration.
For a CNF formula over natural-number variables, its fragment is the bit string obtained by emitting, for each clause in order, a leading clause-marker bit followed by the serialization of that clause; no formula terminator is appended.
For every input word, input position, stream prefix, field word \(w\), stream tail, and old target, there is a strictly positive time, at most twice the old target length plus \(3|w| + 6\), at which the unrestricted loader first reaches its completed control, holding exactly \(w\) with restored target head and advanced stream cursor; no bound relating the old and new lengths is assumed.
For every input word, input position, stream prefix \(u\), field word \(w\), stream tail, and old target, there is a time \(t \le 2\, (\text{old length}) + 3|w| + 6\) after which the unrestricted loader, started at the reset phase with stream cursor \(|u|\), holds exactly \(w\) on its target with head zero and stream cursor just past the consumed field; the native input and physical output are unchanged, and the stream moves only during reading.
A two-tape machine first runs the scratch reset on its target, then dispatches once into the banked field reader on the same tapes: the old target is actually cleared before parsing, so the new field may be empty, shorter, or unrelated to the old contents. Stream positions and all reset and read transitions are charged.
For a template kind together with an input-variable count \(m\), a source time, a target time, and a selected input index, the clause group is the fixed CNF table of that kind’s reconstruction predicate, relabeled along the window wiring into global variable indices.
Given an input length \(n\), a tape count \(k\), a horizon \(T\), and six clause-family generators, the resulting list arranges the clause groups in a fixed order: \(n\) pinning groups, one initial group, \(T\) state groups, \(T+1\) input groups, \(k(T+1)\) work groups ordered by time and then by tape, and \(T\) acceptance groups. An empty clause group still occupies its own slot.
For every list \(L\) of CNF formulas, the list obtained by looking up the entries of \(L\) at indices \(0,\dots ,|L|-1\), with the empty formula as default for out-of-range lookups, is \(L\) itself; the default is never used at a legal index.
For every input length \(n\), tape count \(k\), horizon \(T\), and any choice of the six clause-family generators, the ordered list of clause groups has length exactly \(n + (k+3)T + k + 2\), one more than the last round index; groups are counted, not individual clauses.
For all parameters \(n\), \(k\), \(T\) and all six clause-family generators, concatenating the round chunks of the six-family group list over rounds \(0,\dots ,R\), where \(R = n+(k+3)T+k+1\) is its computed last round, equals the serialization of the flattened six-family formula, without rearranging any clause.
From an encoded preparation header, the full layout places blank counters, an empty record, the original fifth tail as the unary clock, and the fourth field as the virtual reference input into the recorder bank, followed by the four retained arithmetic and instance fields in a disjoint bank; the clock is the original horizon, not a guessed or enlarged one.
For every source machine, parameters \(C, e, c, A, d\), and instance \(x\), with \(Q = C(|x|+1)^e\), \(m = |x| + Q\), and \(T = c(A(m+1)^d+1)^2\), the recorder layout of the arithmetic header equals the recorder’s prepared words over the all-false reference input of length \(m\) at horizon \(T\), extended by the four retained fields: the instance \(x\) and the binary representations of \(Q\), \(m\), and \(T\). The certificate length stays exact.
For every source machine, the function sending an encoded header to the packed self-delimiting encoding of its full recorder layout is polynomial-time computable: every field is a constant empty word or a fixed composition of header projections, and the number of fields depends only on the machine.
For every input word \(x\), copied word \(w\), and prefix length \(j \le |x|\), running the copier in its rewind control with the input head at \(j\) and work head at \(j - 1\) for \(j + 1\) steps returns both heads to their origins and enters the return control; the physical left boundary identifies the origin without an unbounded control counter.
For every input word \(x\), there is a time \(t\) with \(0 \lt t \le 2|x| + 2\) at which the input copier, started from its native initial configuration, first reaches its return control, holding exactly \(x\) on its work tape with both heads restored.
For every input word \(x\) and every \(j \le |x|\), running the input copier for \(j\) steps from its start configuration copies exactly the first \(j\) bits of \(x\) onto the work buffer, with the input head at position \(j + 1\) and the work head at \(j\); every source read and target write is charged.
For every input word \(x\), running the input copier from its native initial configuration for \(2|x| + 2\) steps reaches the configuration assembled from words in which the single work tape holds exactly \(x\), both heads are restored, and the physical output is empty.
For an input-variable count \(m\) and a time \(t\), if the oblivious input head at time \(t\) sits at an interior position \(p\) with \(0 \lt p \le m\), the group is the present-input table wired to input variable \(p-1\); otherwise it is the blank-input table with a safe dummy index of zero.
For an input head position \(p \in \{ 0, 1, \dots , n+1\} \) and a direction \(d\), the clipped direction is stationary when \(p = 0\) and \(d\) is leftward, or when \(p = n + 1\) and \(d\) is rightward, and equals \(d\) in every other case. It expresses native input clamping as a virtual tape-shift direction.
For every input head position \(p \in \{ 0, 1, \dots , n+1\} \) and every direction \(d\), the position reached by the native clamped input move of \(p\) in direction \(d\), shifted down by one, equals \(p - 1\) plus the clipped direction viewed as an integer.
The function taking an ordered flag word to the code of its last marked index is polynomial-time computable: the flag word is threaded through the last-true stripper, an absent result maps to the empty code, and otherwise the stripped prefix’s binary length is exactly the index of the last marked row.
For every oblivious machine, reference length \(m\), time \(t\), and work tape \(\tau \), the last-visit code of the flag word testing, for each \(s \lt t\), whether the work position at \(s\) equals the position at \(t\), is exactly the optional-time code of the public last-visit value; the time is in canonical binary.
Suppose a host machine agrees, on every embedded non-exit control, with the disjoint-bank lift of a smaller machine acting on the left tapes. Then for every start configuration of the smaller machine, every retained right-bank contents and heads, and every horizon \(t\) such that the smaller machine does not reach the exit control before \(t\), the host run agrees with the lifted run of the smaller machine at every time up to \(t\). No agreement is required after the first exit.
For every input word, input position, old fields of length at most \(U\), new fields of length at most \(W\), and stream prefix and trailer, there is a time \(t \le l\, (2U + 3W + 7)\) after which the row loader holds exactly the new row on its field tapes, with every field head restored, the full record stream retained, and the stream cursor immediately after the packed row; the load is a native call, not indexed list access.
For every input word, input position, target field \(i\), old field words, stream prefix \(u\), field word \(w\), and stream tail, there is a strictly positive time \(t\), at most twice the old field’s length plus \(3|w| + 7\), after which the row loader has replaced field \(i\) by \(w\), advanced its field index to \(i + 1\), and moved the stream cursor exactly past the consumed field, from \(|u|\) to \(|u| + 2|w| + 2\); the cost includes target reset, native parsing, target rewind, and one dispatch.
For a positive field count, every input word, input position, old fields of length at most \(U\), new fields of length at most \(W\), and stream prefix and trailer, there is a strictly positive time \(t \le l\, (2U + 3W + 7)\) at which the row loader first reaches its complete-row return, holding the new row with all heads restored and the exact sequential stream cursor.
For every input word, input position, target field \(i\), loader control, old field words, stream, replacement word, and cursor, relocating the two-tape loader configuration into the row loader’s slots, retaining all other fields, yields the row seam in which only field \(i\) has been replaced; every other field remains untouched.
For every input word, input position, old fields of length at most \(U\), new fields of length at most \(W\), stream prefix and trailer, and count \(j \le l\), there is a time \(t \le j\, (2U + 3W + 7)\) after which the row loader, started at field zero with the stream holding the packed new row, has loaded exactly the first \(j\) new fields, restored every field head to zero, and advanced the stream cursor by the exact length of the packed first \(j\) fields. Every field is reset even when revisiting a short row after a longer one.
For every row of field words, coordinate \(i\), stream prefix, and trailer, the stream consisting of the prefix, the packed encoding of the whole row, and the trailer equals the prefix extended by the packed first \(i\) fields, followed by the self-delimiting encoding of field \(i\), whose payload is the packed remaining fields and the trailer. The suffix is preserved literally.
For an input word, an input position, a field index, a loader control, field words, a stream, and a stream cursor, the row-loader seam holds every field as a buffer at head zero and the stream with its explicit cursor; the physical output is empty.
For \(l\) field tapes plus one shared stream tape, the sequential loader is a controller that loads a fixed row in increasing field order: for each field it runs the unrestricted loader, reset followed by parse, on that field and the stream, dispatching once after each completed field; every field clears its own target first, so repeated queries need no monotone-width assumption.
Suppose a host machine with the same tapes agrees, on every embedded control satisfying a guard, with a smaller machine’s transition after relabeling states. Then for every start configuration and horizon \(t\) whose strictly earlier controls all satisfy the guard, the host run equals the state-relabeled run of the smaller machine; all tape data is preserved exactly.
For distinct positive and negative coordinates, every comparator control, candidate row, target words, stream, cursor, clock data, and flags, relocating the four-word comparator configuration into the search’s slots yields the search seam in the comparison phase; the remaining row fields, stream cursor, clock, and accumulated flags are restored literally from the frame.
For every verdict bit \(b\) and search seam in the comparison-finished control, one matcher step appends \(b\) to the flag tape by one real write, advances the query clock by one token, and returns to the clock-test phase; the physical output stays silent.
For distinct positive and negative coordinates, a candidate row and target words of length at most \(W\), and any stream, cursor, clock data, and flags, there is a time \(t \le 2W + 3\) and a bit \(b\), true exactly when the row’s positive count plus the second target equals the first target plus the row’s negative count, such that the matcher completes the comparison, restores all four word heads, stores \(b\) on the flag tape, and advances the clock; cross-sum equality, not the raw words, determines the flag.
For distinct coordinates, a row family with fields of length at most \(W\) among the first \(N\) rows, target words of length at most \(W\), and a trailer, there is a time \(t \le N(l(5W+7) + 2W + 5) + 1\) after which the matcher is in its completed control holding exactly the \(N\) comparison flags in source-time order; the next row and the arbitrary tail are untouched, and at \(N = 0\) only the silent stop executes.
For a positive field count, distinct coordinates, old and new candidate rows of field length at most \(W\), target words, and a stream exposing the packed new row after a prefix, there is a time \(t \le l(5W+7) + 1\) after which the matcher has loaded the new row, cleared all old field targets, advanced the stream cursor exactly past the packed row, and dispatched to the signed comparison with fresh comparator registers.
For every loader control, old and new candidate rows, target words, stream, cursors, clock data, and flags, relocating the row-loader configuration into the search’s slots, retaining the search’s own administrative tapes, yields the search seam in the loading phase; both target words, the clock, and all stored flags are framed, while the candidate row and stream cursor may change.
For distinct coordinates, a row family with all fields of length at most \(W\) among the first \(N\) rows, target words of length at most \(W\), a trailer, and every \(j \le N\), there is a time \(t \le j\, (l(5W+7) + 2W + 5)\) after which the matcher, started with empty candidates, cursor zero, and no flags, holds exactly the comparison flags of rows \(0,\dots ,j-1\) in order, with the stream cursor at the end of the first \(j\) rows and the immutable target words framed throughout.
For distinct coordinates, old and new candidate rows and target words of field length at most \(W\), a stream exposing the packed new row, and a clock position \(j \lt N\), there is a time \(t \le l(5W+7) + 2W + 5\) after which the matcher has consumed one clock token, loaded row \(j\), compared cross-sums, appended exactly the comparison flag of that row, and advanced its cursor past the row; every reset, stream read, comparison, dispatch, and flag write is included in the bound.
For a matcher control, candidate fields, two immutable target count words, the record stream with an explicit cursor, a unary query clock of \(N\) tokens with its head at \(j\), and the stored flag prefix, the search configuration holds every binary field at head zero, the flag tape head at the flag count, and empty physical output.
For distinct positive and negative counter coordinates, the four comparator tapes are embedded as: the candidate’s positive count field, the second target slot, the first target slot, and the candidate’s negative count field; the injection is fixed, with no data-dependent or uncharged record lookup.
For every oblivious machine, reference length \(m\), times \(s\) and \(t\), and work tape \(\tau \), the comparison flag of the stored row at time \(s\) against the target counts at time \(t\), taken at the positive and negative slots of \(\tau \), is true exactly when the public work-head position on \(\tau \) at time \(s\) equals the position at time \(t\).
For distinct positive and negative counter coordinates, the matcher repeatedly tests its unary query clock, loads one stored row from the record stream, compares cross-sums against two retained target words, and appends one flag per row in source-time order; when the clock is exhausted it stops before inspecting the next row, so exactly the clock’s length of rows is inspected and the row at the target time is excluded.
For positive and negative counter coordinates and a stored row, together with two retained target count words, the four comparator words are, in order: the row’s positive count, the second target word, the first target word, and the row’s negative count, exactly the cross-sum interface of the banked comparator.
For a source state, a position tag bit, a scanned virtual input symbol, and scanned work symbols, the displacement vector records, per coordinate, the clamped virtual input displacement and each work tape’s signed move from the source transition; a halted source contributes zero on every coordinate. The input component is the actual movement, not the possibly blocked requested move.
Let \(c\) be a source configuration with a valid position tag bit. For every coordinate, the reference position of the stepped configuration equals the reference position of \(c\) plus the recorded displacement, including blocked input moves and the halting transition.
Let \(c\) be a source configuration and let \(b\) be any valid position tag for its input position. The displacement vector computed from \(b\) equals the displacement vector computed from the deterministic nonzero-position tag; at an interior cell either tag yields the same clamped movement.
For every polynomial-time computable string function \(f\), there are a machine with at least one work tape, entry and exit controls, and constants \(a, d\) such that for every physical input and argument word, the machine started at the entry over the argument’s state word reaches, at a strictly positive first exit within \(a(|arg|+1)^d\) steps, exactly the state word of \(f\) of the argument, with all scratch cells and heads restored after each request.
If every function in a finite list of string functions is polynomial-time computable, then so is the function sending an input \(x\) to the packed self-delimiting encoding of the list of their values at \(x\); all repeated input scans are charged by composition.
For every fixed bit \(b\), the function sending an input \(x\) to the word of \(|x|\) copies of \(b\) is polynomial-time computable, witnessed by an actual native scan of the input rather than by substituting the instance for the simulator’s physical input.
Let \(f\) be polynomial-time computable, and let a machine compute \(g(x)\) from \(f(x)\) within \(b\, (|f(x)|+1)^r\) steps for every input \(x\), correctness being required only on the image of \(f\). Then \(g\) is polynomial-time computable; the produced intermediate word is bounded by the first machine’s runtime and the full buffered-composition ledger is charged.
If \(g\) is a polynomial-time computable string function, then so is the function that decodes its input as an encoded pair \((a,b)\) and returns the encoded pair \((a, g(b))\), returning the empty string when the input is not an encoded pair. The first field is retained with its exact bytes.
For all constants \(C\) and \(e\), the function sending a string \(x\) to the word of exactly \(C(|x|+1)^e\) true bits is polynomial-time computable, including the zero-coefficient and zero-degree cases; the exact value is never replaced by a majorant.
Let \(L\) be a language in \(\mathrm{NP}\). Then there exist certificate parameters \(C\) and \(e\), a verifier language \(V\), a finite oblivious machine \(M\), and constants \(c, A, d \gt 0\) such that a string \(x\) belongs to \(L\) exactly when some certificate \(u\) of length \(C(|x|+1)^e\) makes the concatenation of \(x\) and \(u\) belong to \(V\), and \(M\) decides \(V\) within time \(c\, (A(n+1)^d+1)^2\) on inputs of length \(n\). Neither certificate parameter is enlarged by the normalization.
Let \(V\) be a language in \(\mathrm{P}\). Then there exist a finite machine \(M\) and constants \(c, A, d \gt 0\) such that \(M\) is oblivious, meaning its head positions depend only on the input length and the time, and \(M\) decides \(V\) within time \(c\, (A(n+1)^d + 1)^2\) on inputs of length \(n\). Only the time bound is enlarged; the language, and hence the certificate relation, is unchanged.
Let \(W\) be a binary machine with \(k\) work tapes, \(a\) a padding constant, and \(c\) a configuration of \(W\) on an input \(x\). The clock-stage configuration of the schedule controller on the block-encoded input has control state the clock-phase encoding of the state of \(c\) (a designated reset phase once \(c\) has halted), input head at the position of the input head of \(c\), the \(k\) leading work tapes carrying the bit-encoded work tapes of \(c\), the budget tape capturing the entire append-only output of \(c\) with head one past the captured word, the unary tape holding only an origin sentinel with head at coordinate one, a blank guide tape with head at the origin, and empty output.
Let \(W\) be a binary machine, let \(a\) and \(b\) be constants, and let a time bound \(T\) be such that on every input \(x\), the machine \(W\) outputs the binary representation of \(T(|x|)\) within \(b\, (T(|x|)+1)\) steps. Then for every input \(x\) there exist a time \(\tau \le b\, (T(|x|)+1)\) and a halted configuration \(c\) of \(W\) on \(x\) whose output is the binary representation of \(T(|x|)\), such that running the schedule controller for \(\tau + 1\) steps from its initial configuration on the block-encoded input equals the clock-stage configuration of \(c\). No assumption that \(W\) is oblivious is used.
Let \(W\) be a binary machine, \(a\) a padding constant, and \(x\) a binary input. One step of the schedule controller from its initial configuration on the block-encoded input equals the clock-stage configuration of the initial configuration on \(x\) of the masked clock machine obtained from \(W\); this single transition installs both sentinels and starts the clock.
Let \(W\) be a binary machine, \(a\) a padding constant, and \(c\) a configuration of \(W\) on an input \(x\). After forgetting the bit values (retaining only whether a symbol is present), the symbol read by the input head of the clock-stage configuration of \(c\) agrees with the correspondingly masked native read of \(c\).
Let \(W\) be a binary machine, \(a\) a padding constant, \(x\) a binary input, and \(\tau \) a time bound such that the masked clock machine obtained from \(W\), started on \(x\), is still live at every time strictly before \(\tau \). Then for every \(t \le \tau \), running the schedule controller for \(t + 1\) steps from its initial configuration on the block-encoded input equals the clock-stage configuration of the masked clock run for \(t\) steps from its initial configuration on \(x\); every emitted budget bit is thereby represented.
Let \(c\) be a configuration of the \(k\)-tape Boolean machine \(W\) on a Boolean input that has halted (its state is the halting state). Then the clock-stage embedding of \(c\) into the oblivious schedule machine — which keeps the work block of \(c\) and captures its entire output word on the budget tape — is itself a setup configuration: base the embedded configuration itself, phase reset-start, the embedded input-head position, budget head at cell \(m+1\) over the captured clock tape of the output word of \(c\) (where \(m\) is that word’s length), unary head at cell \(1\) over the unary tape with zero markers, and guide head at cell \(0\) over the blank tape.
Let \(W\) be a binary machine, \(a\) a padding constant, and \(c\) a live configuration of \(W\) on an input \(x\) (its control state is some state \(q\)). Then one step of the schedule controller from the clock-stage configuration of \(c\) equals the clock-stage configuration of one step, from \(c\), of the masked clock machine obtained from \(W\) by ignoring the input bit values; in particular any bit emitted by that step, including a final emission, is captured on the budget tape, and administrative work begins only after this transition.
Let \(w\) be a bit word and consider its captured clock tape: the tape with the origin sentinel at cell \(0\), the \(i\)-th bit of \(w\) at cell \(i+1\) for \(0 \le i \lt |w|\), and blanks elsewhere. A cell of this tape carries the origin mark if and only if it is cell \(0\).
For every bit word \(w\), the captured clock tape of \(w\) (origin sentinel at cell \(0\), the bits of \(w\) on cells \(1,\dots ,|w|\), blanks elsewhere) equals the zipper tape with frontier at cell \(1\), left list consisting of the single origin sentinel, and right list the bits of \(w\), each encoded as a budget-bit symbol.
Let a machine reach, in \(a\) steps from its native start, a configuration in an absorbing completion control with empty output, whose chosen tape holds exactly the word \(w\) as a buffer with its head at the right blank, and whose start control differs from the completion. Then the replay-extended machine computes \(w\) from the same input within \(a + 2|w| + 4\) steps, with an explicit replay ledger and every other tape framed.
Let a machine reach, in \(a\) steps from its native start, an absorbing completion whose chosen tape holds the word \(w\) with its head at some position \(j \le |w|\), with empty output and a start control distinct from the completion. Then the replay-extended machine computes \(w\) from the same input within \(a + j + |w| + 4\) steps; the whole source frame is preserved and every rewind and output transition appears in the ledger.
Given a machine with an absorbing live completion control and a chosen buffer tape, the extended machine runs the original computation unchanged, dispatches once at the first completion, and then replays the chosen buffer to the physical output; the source computation and its startup are untouched before the return.
Given a count \(m\) of input variables, a block width \(B\), a time \(t\), and a bit position \(i\), the packed global variable index is \(m + tB + i\): snapshot bits occupy consecutive blocks of width \(B\) placed after all \(m\) input variables.
For a source machine, parameters \(C, e, c, A, d\), and an instance \(x\), the packed record is the nested pair of: the untouched exact arithmetic header of \(x\); the inclusive chronological movement-count trajectory over the all-false reference input of length \(|x| + C(|x|+1)^e\) at the producer horizon plus one; and the inclusive time-ordered, work-tape-ordered table of canonical last-visit codes.
For every source machine and parameters, there are a machine and constants \(a, r\) such that the machine computes the complete packed preparation data within \(a(n+1)^r\) steps on inputs of length \(n\), and the packed answer itself has length at most the same bound.
For every source machine and parameters \(C, e, c, A, d\), the function sending an instance to its complete packed preparation data, header, trajectory, and visit table, is polynomial-time computable; every retained copy and capture is charged by the native composition calculus.
For a machine with a state code and each template kind, a Boolean predicate on windows of \(2B+1\) bits, where \(B\) is the block width: the initial kind demands that the target block encode the start state reading the selected input bit, or a blank when no input is present; the state kind, that the target state slice encode the stepped state of the decoded source block; the input kind, that the target input slice encode the selected bit or a blank; the work kind for tape \(\tau \), that the target work slice at \(\tau \) encode the symbol written or kept at the last earlier visit, or a blank on a first visit; and the acceptance kind, that the decoded source snapshot does not emit the output bit false. All kinds share the same constant arity, and unused window bits are harmless.
Let a row of field words of length at most \(W\) and an arbitrary trailer be given, and let the native input be the packed encoding of the row followed by the trailer. Then there is a time \(t \le 2|x| + 3 + l(3W + 7)\), where \(x\) is that input, after which the preparation machine holds every parsed field at head zero, retains the original encoded input on its stream tape, and sits with the cursor immediately before the trailer.
For every row of field words of length at most \(W\) and trailer, there is a strictly positive time, within the preparation ledger \(2|x| + 3 + l(3W+7)\), at which the preparation machine first reaches the prepared seam with all fields parsed; native input copying and the dispatch ensure positivity even with zero fields.
For every field count \(l\) and input word \(x\), there is a time \(t \le 2|x| + 3\) after which the preparation machine, started from its own native initial configuration with no assumed work-tape words, sits at the row loader’s entry: all field targets blank, the stream holding exactly \(x\) with cursor zero, and both heads canonical.
Let work words of length at most \(W\) and a trailer be given, and let the native input be their packed encoding followed by the trailer. Then there is a startup time \(a \le 2|x| + 4 + k(3W+7)\), where \(x\) is that input and \(k\) the field count, such that for every duration \(b\), running the loading-extended machine for \(a + b\) steps equals the frame around the supplied machine’s \(b\)-step run from the configuration assembled over the parsed words; no cleanup or random access is assumed.
For a machine expecting prepared work words, the extended machine first parses arbitrary exact work words from its self-delimiting native input into the field bank, keeping the retained input on its own extra tape, then dispatches once and runs the supplied machine on the loaded bank.
For \(l\) field tapes plus one stream tape, the preparation machine first runs the native input copier into the stream slot, then dispatches once into the sequential row loader, which parses the fixed number of self-delimiting fields from the copied stream.
For parameters \(C, e, c, A, d\) and an instance \(x\), set \(Q = C(|x|+1)^e\), \(m = |x| + Q\), and \(T = c\, (A(m+1)^d+1)^2\). The header is the nested pair encoding of six fields: the instance \(x\), the binary representations of \(Q\), of \(m\), and of \(T\), the all-false reference input of length \(m\), and a unary clock of \(T\) true bits. Trajectory and last-visit records are not part of it.
For all parameters \(C, e, c, A, d\), the arithmetic preparation header, comprising the retained instance, the exact binary certificate length, reference length, and horizon, the all-false reference input, and the unary horizon clock, is computable in polynomial time by native machines with sequential access; no positivity assumption is placed on any parameter.
For a source machine, work tape \(\tau \), and polynomial-time producers of a reference length (as an all-false word) and a unary time, the function sending an input to the optional-time code of the public greatest strictly earlier visit to the currently scanned cell of \(\tau \) is polynomial-time computable, with the canonical binary time encoding.
Suppose the last-visit function for input count \(m\), time \(t\), and tape \(\tau \) returns a time \(s\). Then \(s \lt t\), the work-head position on \(\tau \) at time \(s\) equals the position at time \(t\), and every time \(r \lt t\) scanning the same position satisfies \(r \le s\).
For all parameters \(C, e, c, A, d\), the function sending an input \(x\) to the unary word of the producer horizon at \(|x|\) plus one true token is polynomial-time computable; the horizon is read off the inherited arithmetic header, and the extra token accounts for the inclusive final row.
For all naturals with \(N \le L\), the complete query ledger, twice the scan cost \(2L + 9 + (l+5)(3L+7) + N(l(5L+7) + 2L + 7)\) plus \(c(N+1)^e + 2\), is at most \((100(l+10) + c)(L+1)^{e+2}\), one common monomial in the encoded argument length \(L\).
For all coordinates, input word, target words, stream, and bound \(N\), the configuration assembled from the query argument words at the matcher’s start control equals the search seam with empty candidates, cursor zero, full clock, and no flags, including every scratch tape and head.
For distinct positive and negative coordinates, there exist a single machine \(Q\) and constants \(c, e\) such that on every canonical query argument, a packed row family, target words, and strict-prefix clock, \(Q\) computes the optional-time code of the last row whose cross-sums match the targets, within twice the query ledger plus \(c(N+1)^e + 2\) steps. Only canonical inputs are promised; callers must establish them.
For distinct coordinates, a row family with fields of length at most \(W\) among the first \(N\) rows, target words of length at most \(W\), a trailer, and argument words of length at most \(U\), the query transducer computes, from the packed encoding of the query argument, exactly the list of the \(N\) comparison flags in source-time order, within \(2|a| + 9 + (l+5)(3U+7) + N(l(5W+7) + 2W + 7)\) steps, where \(a\) is the argument; argument copying, preparation, every row query, comparison, dispatch, flag write, rewind, replay, and the halt are all charged.
For distinct positive and negative coordinates, the query transducer prepares its field bank from the actual encoded input, runs the all-earlier-row matcher over exactly the clocked strict prefix, and then physically replays the stored flag word as its output.
For a field count \(l\), two target count words, a record stream, and a strict-prefix bound \(N\), the query argument words are: blank candidate fields, the retained stream, the two target counts, a unary clock of \(N\) tokens, and a blank flag tape.
For every input word, input position, field word \(w\), stream prefix, stream tail, written target prefix, and old suffix, running the reader for \(2|w|\) steps decodes all of \(w\) into the target, one overwritten cell per decoded bit, preserves the exact unread stream suffix, and leaves the separator for the following steps; both stream bits of each pair are charged individually.
For every input word, input position, stream consisting of a prefix \(u\) followed by a doubled bit and a remainder, and target split as a written prefix and an old suffix, two reader steps from the scanning control consume the two equal stream bits, overwrite the first old target cell with the decoded bit, and advance the stream cursor by two and the target cursor by one, using real adjacent head moves.
For every input word, input position, stream, decoded target \(w\), stream cursor, and count \(j \le |w|\), running the reader in its rewind control at target cell \(j-1\) for \(j+1\) steps returns the target head to zero and enters the completed control; the stream cursor never moves during the rewind.
For every input word, input position, stream prefix \(u\), field word \(w\), stream tail, and old target of length at most \(|w|\), running the reader for \(3|w| + 3\) steps from the scanning control at stream cursor \(|u|\) yields the completed control, in which exactly one encoded field is consumed, the target is \(w\) with its head restored to zero, the whole stream is preserved, and the stream cursor sits just past the consumed field; the empty field takes three steps.
For every input word, input position, stream prefix followed by the false-true separator and a tail, and target word, two reader steps yield the rewinding control, with the separator consumed, the stream cursor advanced past it to the next field, and the target head moved once left, without writing the separator and without changing the stream or target words.
For an input word, an input position, a reader control, a stream word, a target word, and independent stream and target cursors, the reader configuration holds the stream and target as buffers with the given heads; physical input and output are fixed.
A machine with a read-only stream tape and a target tape consumes one aligned doubled-bit field: it reads each stream bit into finite control, checks the repeated bit, overwrites one target cell per decoded bit, detects the false-true separator, rewinds only the target head, and idles in a completed control. Its contract assumes the old target is no longer than the decoded field, as holds for movement counts read in increasing source-time order.
Let \(c\) be a source configuration with a valid position tag, and let the movement counts have binary representations of width at most \(W\). From the source-release phase, there is a time \(s \le 2W + 5\) after which the recorder has executed exactly one source transition together with its full signed counter update and dispatched to the next row-copy phase, with a valid tag for the stepped configuration; the record and clock contents and heads are framed throughout.
For every source machine, physical input, virtual input \(y\), and horizon \(T\), there is a time \(s \le (T+1)(F(3T+4) + 2T + 7)\), where \(F\) is the field count, after which the recorder, started from the prepared blank seam, is in its completed live control with the stored rows of all times \(0,\dots ,T\) on the record tape, the movement counts and the exact horizon configuration of the source, and an empty physical output.
For every source configuration, tag, movement counts with binary representations of width at most \(W\), record, clock budget, and logical time, there is a time \(s \le F(3W+4) + 1\), where \(F\) is the field count, after which the recorder has appended the complete stored row of the counters to the record and dispatched into its clock-test phase; all source storage and the logical clock are unchanged throughout.
For every source machine and horizon \(T\), the complete stored-row ledger \((T+1)(F(3T+4) + 2T + 7)\), where \(F\) is the field count, is at most \((4F + 7)(T+1)^2\), so native recording work is polynomial in the source horizon, separately from the formula output size.
For every source machine, physical input, virtual input \(y\), horizon \(T\), and logical time \(j \le T\), there is a time \(s \le j\, (F(3T+4) + 2T + 7)\), where \(F\) is the field count, after which the recorder, started from the prepared blank seam, holds on its record tape exactly the stored rows of times \(0,\dots ,j-1\), carries the movement counts and source configuration of time \(j\) with a valid tag, and is ready to store the current row. The invariant concerns actual tape contents, not merely a running schedule.
For every source machine, physical input, virtual input \(y\), and horizon \(T\), the configuration assembled from words, blank movement counters, empty record, a unary clock of \(T\) true cells, the virtual input, and blank source work tapes, equals the recorder’s prepared seam at logical time zero over the initial source configuration. No claim is made about native initialization from raw input.
For every source configuration, tag, movement counts, record, clock budget, logical time, field index, and copier control, relocating the row-copy seam into the recorder’s tapes, retaining every inactive slot from the recorder configuration itself, yields exactly the recorder configuration in the row-copying phase; the inactive bank includes every source cell and head and the unchanged clock.
For every source configuration, tag, movement counts, clock budget, and logical time, the retained inactive tapes and heads used when relocating a row-copy configuration into the recorder are the same for any two record contents and row controls, so consecutive fields and rows can append to the record without changing source time.
For every source configuration, tag, movement counts, record, and clock budget \(T\), one recorder step from the clock-test phase at logical time \(T\) reads the exhausted clock blank and yields the completed live control, silently; no further source transition can occur.
For every source configuration, tag, movement counts, record, clock budget \(T\), and logical time \(t \lt T\), one recorder step from the clock-test phase reads the live unary clock cell, advances only the clock head from \(t\) to \(t+1\), and enters the source-release phase; no source head moves and no record bit is appended.
For every source configuration, tag, movement counts, record, clock budget, and logical time, relocating the canonical tracking seam into the recorder’s tapes, retaining the record and clock slots from the recorder configuration itself, yields exactly the recorder configuration in the counter-update phase.
The retained inactive tapes and heads used when relocating a tracking configuration into the recorder depend only on the stored record and clock, so they agree for any two source configurations, tags, and movement counts; source writes and counter updates therefore do not disturb the frame.
For a recorder control, a source configuration, movement counts, a stored record, a clock budget \(T\), and a logical time \(t\), the recorder configuration holds every movement counter in canonical binary at head zero, the actual record prefix with its head at the record length, the unary clock with its head at \(t\), and the complete virtual reference bank of buffer and source tapes; only the record and clock heads may be away from zero.
Let a source machine, virtual input \(y\), horizon \(T\), and four retained fields be given, with every laid-out field of length at most \(W\), and let the native input be the packed encoding of the layout. Then within \(2|x| + 4 + (k'+4)(3W+7) + (4F+7)(T+1)^2\) steps, where \(x\) is that input, \(k'\) is the recorder tape count, and \(F\) the field count, the host reaches the completed recorder configuration holding all stored rows of times \(0,\dots ,T\), the horizon source configuration, and the protected header fields and input buffer.
For every source machine, input word, virtual input \(y\), horizon \(T\), and retained fields, the loader’s completed configuration over the parsed recorder words, with its control replaced by the recorder’s start, equals the host configuration around the recorder’s prepared seam at logical time zero; source initialization, counters, clock, and record bank are all fixed by the established prepared-seam identity.
For all naturals, if a recording time is at most \(2n + 4 + k(3W+7) + (4l+7)(T+1)^2\) and a replayed length is at most \(2l(T+1)^2\), then the recording time plus twice the replayed length plus four is at most \(2n + 8 + k(3W+7) + (8l+7)(T+1)^2\).
For an input word, four retained header fields, and a recorder configuration, the host configuration is the recorder configuration embedded on the active slots, with the four header fields kept at head zero and the original encoded input kept on the final buffer with its cursor at the input’s full length.
For every source machine, virtual input \(y\), horizon \(T\), retained fields, and width bound \(W\) covering all laid-out words, the trajectory-output machine computes, from the packed encoding of the layout, the complete serialized inclusive trajectory of rows \(0,\dots ,T\), within \(2|x| + 8 + (k'+4)(3W+7) + (8F+7)(T+1)^2\) steps, where \(x\) is the input, \(k'\) the recorder tape count, and \(F\) the field count; preparation, recording, rewind, replay, and the terminal transition are all charged.
For every source machine, virtual input \(y\), horizon \(T\), and retained fields, the trajectory-output machine computes the inclusive trajectory from the packed layout within a constant times the square of the encoded argument length plus one; every initial word and the unary horizon occur inside the argument, so the bound covers a zero horizon as well.
For a source machine and a virtual input \(y\), the serialized trajectory prefix at count \(t\) is the concatenation of the stored rows of the movement counts at times \(0, \dots , t-1\), each row consisting of all its self-delimiting binary fields in tape order; the complete inclusive trajectory at horizon \(T\) uses prefix length \(T+1\).
For every source machine, virtual input \(y\), horizon \(T\), and count \(t \le T + 1\), the serialized trajectory prefix of the first \(t\) rows has length at most \(t\) times the field count times \(2T + 2\): every stored count has width at most the horizon and every field is exactly a doubled payload plus its separator.
For every source machine and parameters \(C, e, c, A, d\), the function sending an instance \(x\) to the inclusive serialized trajectory over the all-false reference input of length \(m = |x| + C(|x|+1)^e\) at horizon \(T = c(A(m+1)^d+1)^2\), with \(T+1\) rows, is polynomial-time computable; the output is produced by the actual recorder, not read off a prepared configuration.
The host machine runs native field preparation on the full bank, dispatches once at the prepared seam into the unchanged inclusive trajectory recorder, and then runs the recorder on its active slots while the retained header fields and input buffer stay protected; the completed recorder remains live so that later search and packing phases can be attached.
If a reference-input producer and a unary-horizon producer are polynomial-time computable, then so is the function packing the recorder’s full argument: blank counters, empty record, the produced unary clock, the produced virtual input, blank source tapes, and four empty retained fields, as self-delimiting fields.
The recorder’s finite control is a disjoint sum of five phases: row copying, carrying the source state, tag, and row-controller control; the clock test; the source-step release; the running counter update, carrying a tracking control; and the completed live return.
For a source machine, the recorder cycles through four phases: it copies the current row of movement counters onto the record, tests the unary clock, releases exactly one tracked source step with its counter updates, and returns to row copying; when the clock is exhausted it enters a completed live control. The current row is stored before the clock test, so both the initial and the final rows are present, and every counter and copy transition lies outside source time.
The recorder’s selection maps each counter-bank tape and each source-side tape back to its tracking-machine coordinate and leaves the record and clock slots unselected, so both administrative record tapes are preserved during tracking.
For a source machine, virtual input \(y\), and horizon \(T\), the recorder’s prepared work words are: empty movement counters, an empty record, a unary clock of \(T\) true bits, the virtual input \(y\), and blank source tapes, in the recorder’s established tape layout.
Let \(c\) be a source configuration over a virtual input, with a valid position tag bit \(b\). Applying the embedded reference action to the embedded host configuration of \(c\) yields exactly the embedded configuration of the stepped source configuration, for some valid new tag bit: each administrative tape is updated by its requested write and head move, while the host input position and the physical output are unchanged.
For an embedded source state, a position tag bit, the host work-tape reads, and arbitrary simultaneous operations on the \(l\) administrative tapes, the host action simulates one source step through the input buffer: a live source state applies its transition’s work-tape writes and virtual input move before its successor state is internalized, so erasures and terminal writes take effect; a halted source keeps all source components fixed. In both cases source emission and physical termination are suppressed, and the bank operations are applied.
Given a source machine, a host-state embedding, a source configuration over a virtual input word, a position tag bit, a host input position, an administrative bank of \(l\) extra tapes with their heads, and an output word, the host configuration places the virtual input on a buffer tape next to the source work tapes and the bank. The source state is embedded into a live host state even when halted, the source output is deliberately absent, and the physical output is the separately supplied word.
For a source configuration \(c\), a position tag bit, a clock budget \(T\), and an elapsed count \(t\), the clocked-runner configuration embeds \(c\) over the virtual input buffer, fills the clock tape with \(T\) true cells and places its head at \(t\), and has empty output, so source time is recorded separately from any future administrative cost.
For a source machine, the host machine carries one unary clock tape, one input buffer tape, and one tape per source work tape, and runs the embedded reference simulation while consuming one clock cell per source step. Exhaustion of the clock dispatches to a distinguished live return state; halting of the source alone does not dispatch. Its initial-configuration contract assumes a prepared configuration with the header installed.
For every source configuration, tag bit, and counter value \(n\), there is a strictly positive time \(t\), at most twice the binary length of \(n\) plus two, such that the counter-beside-reference host first reaches its return control exactly at \(t\), at which point the counter tape holds the binary representation of \(n+1\) with its head restored to zero, while the source configuration, tag, and every bank tape and head are untouched, and the represented source clock has not advanced.
For a source machine, the host is a machine that runs the four-state binary increment on its leftmost tape while retaining the entire virtual reference configuration, buffer tape, source tapes, and heads, in a disjoint bank; the finite control stores the source state and the boundary tag throughout the call.
For all naturals with \(N \le L\), the complete repetition ledger, preparation \(2L + 9 + k(3L+7)\), plus \(N\) calls each costing the clean-call polynomial evaluated at the quadratic state bound \(K(L+1)^2\), plus the final replay \(K(L+1)^2\), is at most \((13 + 7k + a(K+1)^d + K)(L+1)^{2d+3}\).
Suppose a clean call runs from the entry over a state word to its exit over a new state word in a strictly positive time \(t\), never showing the exit control strictly before \(t\). Then the repetition controller, started in its begin phase over the same bank, reaches in exactly \(t\) steps the seam at the exit control with the new state word; the first source action is executed before any return test, so an entry equal to the exit is handled, and the unary clock is preserved.
Suppose every iterate \(j \lt N\) of a word function \(f\) has a clean call of positive duration at most \(B\) from the controller’s bank. Then within \(N(B+2) + 1\) steps the controller, started at its test phase over the initial word with an \(N\)-token clock, reaches its completed control holding exactly the \(N\)-fold iterate of \(f\); every call starts from the fully restored seam and the final blank clock executes one silent completion.
For every clean-call machine, controls, input, state word, clock budget, cursor, and inner control, relocating the configuration assembled over the state word into the repetition controller’s slots, with the protected clock retained, yields exactly the repetition seam at that inner control.
Suppose a clean call runs from the entry over a state word to its exit over a new state word in a strictly positive time \(t\), never showing the exit strictly before \(t\), and let the clock cursor be \(j \lt N\). Then the controller, started at its test phase, advances the clock by one token, executes the entire positive clean call, and performs one return dispatch, reaching the test phase over the new state word at cursor \(j + 1\) in exactly \(t + 2\) steps; calls returning empty words are included.
For a controller state, an argument word, a clock budget \(N\), and a logical cursor \(j\), the repetition seam holds the clean argument bank at blank scratch with the state word installed, and the protected unary clock with its head at \(j\); the physical output is empty.
Suppose every argument has a clean call computing one application of a word function \(f\) within \(a(|arg|+1)^d\) steps, and every iterate of \(f\) from the initial word \(s\) up to \(N\) has length at most \(W\). Then the repetition-output machine computes the \(N\)-fold iterate of \(f\) at \(s\) from the packed repetition argument within \(2|x| + 9 + (k+1)(3|x|+7) + N(a(W+1)^d + 2) + W\) steps, where \(x\) is the argument and \(k\) the call machine’s tape count; input parsing, all positive calls, every dispatch, and the final replay are charged.
The repetition-output machine loads its argument bank natively from self-delimiting input, runs the unary-bounded repetition to completion, and then physically replays the completed state word from the first scratch tape as its output.
Given a clean-call machine with entry and exit controls, the repetition controller tests a protected unary clock, and for each live token executes one complete clean call on the shared argument bank, then dispatches back; a fresh begin phase executes the call’s first transition before testing the positive return, so entry and exit controls may coincide. On a blank clock it completes silently.
For every input word, input position, buffer \(w\), output, and count \(j \le |w|\), running the replay machine in its rewind control at cell \(j - 1\) for \(j + 1\) steps restores the buffer head to zero and enters the emitting control; each occupied prefix cell is read once and the empty-buffer case is included.
For every input word, input position, buffer \(w\), position \(j \le |w|\), and output prefix, running the replay machine in its emitting control for \(|w| - j + 1\) steps appends exactly the suffix of \(w\) from position \(j\) to the output and halts on the right blank, without any spurious trailing bit.
For every input word, input position, buffer \(w\), output prefix, and head position \(j \le |w|\), running the replay machine from its start control at cell \(j\) for \(j + |w| + 3\) steps appends exactly \(w\) to the output and physically halts; only the prefix rewind, the full emission, and the halt are charged.
For every input word, input position, buffer \(w\), and output prefix, running the replay machine from its start control at the buffer’s right blank for \(2|w| + 3\) steps appends exactly \(w\) to the output and physically halts; the initial left move, full rewind, every emitted bit, and the halt are all charged.
A three-state machine with one work tape replays a stored buffer to the physical output: started at the buffer’s right blank, it rewinds explicitly to the origin, emits every stored bit left to right, and physically halts on the final blank; no verdict or terminator is added.
For every input word, input position, selected counter \(i\), counter words \(w\), and record, there is a strictly positive time \(t \le 3|w(i)| + 4\) after which the row controller has appended the self-delimiting encoding of \(w(i)\) to the record and advanced its field index from \(i\) to \(i+1\), with the entire counter bank, the input head, and the empty physical output restored.
For a positive number \(l\) of counters, every input word, input position, counter words of length at most \(W\), and record, there is a time \(t\) with \(0 \lt t \le l(3W+4)\) such that the row controller never shows the end-of-row control before \(t\), and at \(t\) the record has been extended by the complete stored row with all counter heads restored; positivity comes from the fixed positive field count.
For every input word, input position, selected counter \(i\), copier control, counter words, and record, relocating the two-tape copier configuration for counter \(i\) into the row controller’s slots, with all other counters retained, yields exactly the row seam at field \(i\); every unselected counter retains its whole word, not just its current cell.
For every input word, input position, counter words all of length at most \(W\), record, and count \(j \le l\), there is a time \(t \le j(3W + 4)\) after which the row controller, started at field zero, has appended exactly the stored prefix of the first \(j\) fields to the record and sits at field index \(j\) with every counter and head restored.
For every input word, input position, counter words of length at most \(W\), and record, there is a time \(t \le l(3W+4)\) after which the row controller has appended the complete stored row, all \(l\) self-delimiting fields in tape order, to the record, with every source counter unchanged and its head at zero.
For an input word, an input position, a field index, a copier control, stored counter words, and an accumulated record, the row seam holds every counter as a buffer at head zero and the append-only record with its head at the record’s exact current length; the physical output is empty.
For a selected counter coordinate \(i\) among \(l\) counters, the two tapes of a copy call are embedded as: the source slot is counter tape \(i\), and the record slot is the final extra tape. The index is fixed in finite control for the whole field.
For every family of counter words on \(l\) coordinates and every count \(j\), the stored row prefix is the concatenation of the self-delimiting encodings of the first \(j\) words in increasing tape order; indices beyond \(l\) contribute nothing, and actual controllers only use prefixes up to the field count.
For counter words of length at most \(W\) and every count \(j\), the stored row prefix of the first \(j\) fields has length at most \(j(2W + 2)\): each field costs its doubled bits plus the two separator bits, independently of the native copying runtime.
For every family of rows of \(l\) binary fields indexed by time and every count \(n\), the stream of the first \(n\) rows is the concatenation of the packed self-delimiting encodings of rows \(0, \dots , n-1\) in order, in exactly the banked field format.
For every row family, bound \(N\), time \(j \lt N\), and trailer, there is a remainder word such that the full stream of \(N\) rows plus trailer equals the stream of the first \(j\) rows, followed by the packed encoding of row \(j\), followed by that remainder; obtaining the row itself still requires the native loader.
For \(l\) counter tapes plus one record tape, a controller runs the two-tape field copier on each counter in fixed tape order: its control holds the current field index and the copier control, one silent dispatch is charged after each completed field, and after the last field the controller idles in a live end-of-row control.
For a source machine, distinct positive and negative counter coordinates, and polynomial-time reference-input and unary-time producers, the function sending an input to the last-visit code of the flags comparing each strictly earlier row’s cross-sums against the final targets is polynomial-time computable; the scanner consumes exactly the first \(T\) rows of the inclusive \(T+1\)-row trajectory, and the uniform query envelope charges argument construction, comparisons, reverse search, and output.
For a source machine, virtual input \(y\), time \(t\), and two counter coordinates, the query argument packs the query words over the inclusive trajectory of \(t+1\) rows with a strict candidate clock of \(t\) tokens, so the final row is retained as an unconsumed suffix beyond the scanned prefix.
If a reference-input producer and a unary-time producer are polynomial-time computable, then for any two counter coordinates the function sending an input to its complete packed query argument is polynomial-time computable; each field comes from the banked recorder, a native final-count output, the unary clock, or a native constant, and the field packer charges every retained copy.
For a source machine, virtual input \(y\), time \(t\), and positive and negative counter coordinates, the two target fields are the binary representations of the two selected movement counts at time \(t\), taken from the final row of an actual source run; they are compared by cross-sums, never by raw word equality.
Given a displacement vector on \(l\) coordinates, the selection on \(2l\) counters is as follows: for each coordinate, the counter in the first bank is marked exactly on a positive move and the counter in the second bank exactly on a negative move; a zero move selects neither.
For movement-count vectors \(n\) and \(m\) and a coordinate \(i\), the signed positions at \(i\) agree exactly when the two cross-sums of first-bank and second-bank counts are equal; the criterion compares values rather than the non-unique two-counter codes.
For every partial selection, state embedding, retained tapes and heads, action, and local configuration, applying the relocated action to the relocated configuration equals relocating the result of applying the action locally; writes, head motion, and final emission are preserved, and inactive host tapes and heads are fixed.
Suppose a host control acts, on all reads, as the relocation of a smaller machine’s transition at a fixed local state \(q\), through a tape injection whose selection is a left inverse. Then one host step from the relocated configuration of any local configuration in state \(q\), with its control replaced by that release control, equals the relocation of one local step. Executing the first action before observing any return guard handles equal entry and exit controls.
Let a host machine agree, on every embedded control satisfying a guard, with the relocation of a smaller machine’s transition through a tape injection whose selection is a left inverse. Then for every retained host contents and heads, every local start configuration, and every horizon \(t\) such that all local controls strictly before \(t\) satisfy the guard, the host run up to time \(t\) equals the relocation of the local run; the guard is not needed at the endpoint itself.
Given a partial selection from \(l\) host tapes back to \(k\) local tapes and a state embedding, an action on \(k\) tapes is relocated to the host: each selected host tape receives the corresponding local write and move, every unselected host tape is stationary, and the input move, output emission, and embedded successor state are carried over.
Given a partial selection from host tapes to local tapes, a state embedding, and retained contents and heads for the unselected host tapes, a local configuration is relocated to the host: selected tapes and heads come from the local configuration, unselected ones keep their retained values, and the input position and physical output are those of the local configuration.
For every verifier machine with a state code, certificate parameters, instance, and horizon, concatenating the round chunks of the tableau’s ordered group list over all rounds yields exactly the serialization of the tableau formula, including the empty-template and final-terminator cases.
For a verifier machine with a state code, certificate parameters \(C\) and \(e\), an instance \(x\), and a horizon \(T\), the ordered list of clause groups comprises: one unit clause pinning each of the \(|x|\) input variables to the corresponding bit of \(x\), the initial group, state groups for each time below \(T\), input groups for each time up to \(T\), work groups for all times up to \(T\) and all tapes, and acceptance groups for each time below \(T\). The total variable prefix has length \(m = |x| + C(|x|+1)^e\), using the exact original certificate polynomial.
For every arity \(\ell \), predicate \(p\) on \(\ell \)-bit vectors, wiring map \(w\) on variable indices, and assignment \(a\), the relabeling of the chosen CNF table along \(w\) evaluates to \(p\) applied to the bits \(a(w(0)),\dots ,a(w(\ell -1))\); repeated wire indices are permitted and no injectivity is assumed.
For every arity \(\ell \) and predicate \(p\) on \(\ell \)-bit vectors, the chosen CNF table uses at most \(\ell \) variables, has at most \(2^\ell \) clauses, has width at most \(\ell \), and evaluates under any assignment \(a\) to exactly \(p\) applied to the first \(\ell \) bits of \(a\).
The finite catalogue of template kinds for the Cook–Levin tableau over a machine with \(k\) tapes: an initial kind carrying an input-presence flag, a state kind, an input kind carrying a presence flag, a work kind for each tape carrying an earlier-visit flag, and an acceptance kind. Boundary and absent-previous-visit variants are separate fixed tables.
For every source configuration, tag bit, component controls, and counter values, the left-bank lift of the counter-bank product configuration over the retained source tapes coincides literally with the embedded reference configuration whose administrative tapes hold the counters in binary, including the virtual input buffer, every source tape, and all heads.
Let \(c\) be a source configuration with a valid position tag bit, and let movement counts be given whose binary representations fit a common width \(W\). Then there is a strictly positive time \(\tau \le 2W + 4\) such that: every intermediate time strictly between \(0\) and \(\tau \) is in the bank phase, not at the anchor; and at \(\tau \) the tracking host is back at the canonical seam for the stepped source configuration, with a valid new tag and with exactly the clamped displacement counts advanced. Source output is suppressed and every counter head is restored.
Let \(c\) be a source configuration with a valid position tag bit and let movement counts be given in canonical form. One step of the tracking host from the canonical seam applies the full source action, produces a valid tag for the stepped configuration, and enters the bank phase with exactly the counters of the actual displacements selected; no counter has yet changed, and the physical output is still empty.
For a source machine, the tracking host carries two counter banks, one per movement direction, beside the reference bank of buffer and source tapes. In its anchor phase it executes one faithful embedded source action and selects the counters matching the actual clamped displacements; in its bank phase it runs the selected counter increments without advancing the source again, then dispatches back to the anchor. The anchor is a live control phase, distinct from source halting.
If a reference-input producer and a unary-horizon producer are polynomial-time computable, then the function sending an input \(x\) to the inclusive serialized trajectory of the source over the produced reference input, with the produced horizon plus one rows, is polynomial-time computable; no recomputed semantics or unbounded numeric conversion is used.
For every source machine, reference length \(m\), time \(t\), and tape \(\tau \), the optional-time code of the greatest strictly earlier visit has length at most \(t + 2\): a present predecessor is strictly below the query time, and absence at time zero is included.
For a source machine, certificate parameters \(C\) and \(e\), and a request pairing an instance \(x\) with a unary query-time word, the visit row is the packed list, in tape order, of the self-delimiting last-visit codes of every work tape at the query time, over the all-false reference input of length \(|x| + C(|x|+1)^e\).
For every source machine, parameters \(C\) and \(e\), instance \(x\), and time \(t\), the visit row of the encoded pair of \(x\) with \(t\) unary tokens is exactly the packed list, in source tape order, of the optional-time codes of the greatest strictly earlier visits at time \(t\) over the reference length \(|x| + C(|x|+1)^e\).
For every source machine and certificate parameters \(C\) and \(e\), the function sending a retained-instance and unary-time request to the packed row of last-visit codes for all work tapes is polynomial-time computable: native projections recover the instance and time, exact preparation builds the reference input, each tape’s search is native, and finite pairing charges the whole ordered row.
For every source machine and parameters \(C, e, c, A, d\), the function sending an instance \(x\) to the ordered table of visit rows for all times up to the producer horizon plus one is polynomial-time computable: the retained instance and empty table start the loop, a clean call executes each full row step, the repetition controller runs the exact unary number of calls with a quadratic bound on all intermediate state words, and a native projection returns precisely the stored rows.
For every source machine, parameters, instance \(x\) of length at most \(L\), and horizon \(N \le L\), every iterate up to \(N\) of the producer step from the empty retained state has length at most \((8 + 8k)(L+1)^2\), where \(k\) is the number of work tapes.
One internal step maps a retained state to the state with the same instance, the unary clock extended by one token, and the stored rows extended by the complete visit row of the current instance and time; every operation is a native word computation.
For every source machine and parameters \(C\) and \(e\), there are a machine, entry and exit controls, and constants such that each producer step has a clean reusable native call: a strictly positive first exit within a polynomial of the argument length, returning exactly the stepped state word with all scratch restored. The variable number of such calls is not yet assembled here.
For every source machine, parameters, instance \(x\), and count \(t\), iterating the producer step \(t\) times from the empty retained state yields the retained state at time \(t\) holding exactly the ordered table of the first \(t\) visit rows.
For every source machine, parameters, instance \(x\), stored rows, and time \(t\), one producer step on the retained state at time \(t\) yields the retained state at time \(t+1\) whose rows are extended by the visit row of time \(t\) in fixed work-tape order.
For every input word, input position, stream, stream head, and target word \(w\), running the reset machine in its erasing control at cell \(|w| - 1\) for \(|w| + 1\) steps erases every occupied target cell, including the first one, and returns the target head to zero in the completed control; the stream is neither read nor modified.
For every input word, input position, stream, target word \(w\), and stream head, there is a time \(t\) with \(0 \lt t \le 2|w| + 2\) such that the reset machine first reaches its completed control exactly at \(t\), with the target erased, the target head at zero, and the stream untouched.
For every input word, input position, stream, target word \(w\), stream head, and start cell \(j \le |w|\), running the reset machine for \(|w| - j + 1\) steps from its scanning control moves the target head right across each occupied cell, turns at the right blank, and enters the erasing control at cell \(|w| - 1\); the target is still unchanged.
For every input word, input position, stream, target word \(w\), and stream head, running the reset machine for \(2|w| + 2\) steps from the scanning control at target cell zero leaves an empty target with its head at zero, in the completed control; both blank transitions are included and no bound on any later field is required.
A three-state machine with a stream tape and a target tape resets its scratch target: it scans the target to its right blank, erases every occupied cell on the way back to the left blank, and returns at target cell zero in an absorbing control. The stream tape and its cursor are never read or moved.
Given \(m\) input variables, a block width \(B\), a source time, a target time, and a selected input index \(p\), the local window index \(j\) is wired to the \(j\)-th bit of the source block when \(j \lt B\), to the \((j-B)\)-th bit of the target block when \(B \le j \lt 2B\), and to \(p\) otherwise. Repeated global indices are intentional and supported by relabeling.
For an input-variable count \(m\), a time \(t\), and a work tape \(\tau \): when some strictly earlier time \(s\) was the last visit to the cell scanned on \(\tau \) at time \(t\), the group is the earlier-visit table wired from the snapshot block at \(s\); when the cell was never visited before, it is the blank first-visit table.
A total function \(f\) on bit strings (from \(\{ 0,1\} ^*\) to \(\{ 0,1\} ^*\)) is computable if there exists a finite Turing machine over the binary alphabet that computes \(f\): on every input \(x\) it halts, after some finite number of steps, with \(f(x)\) on its output tape. No time constraint is imposed [AB09, §1.4, p. 20].
Let \(N\) be a nondeterministic machine and let the physical input \(y\) decompose as \(y = p\, r\) for words \(p\) and \(r\). From the suffix-copy loader state at position \(|p| + 1\) with buffers \(x\) and \(u\), exactly \(|r|\) steps of the pair-loader machine append \(r\) to the choice buffer, ending at the right boundary of \(y\) with the input buffer \(x\) unchanged and no physical output.
Let \(N\) be a nondeterministic machine over the binary alphabet and let \(d\) be any configuration of the simulation core of \(N\). For every step count \(t\), running the pair-loader machine of \(N\) for \(t\) steps from the state-renamed copy of \(d\) yields exactly the state-renamed copy of the \(t\)-step run of the simulation core from \(d\).
Let \(M\) be a deterministic machine over the binary alphabet that computes an output \(v\) on input \(x\) within \(T\) steps. Then every choice word of length \(T\) drives the guessing-phase machine of \(M\) from its initial configuration on \(x\) to the invariant configuration for the \(T\)-step run of \(M\), with some captured word of length \(|v|\); and conversely, every word \(u\) with \(|u| = |v|\) is the captured word of that configuration for some choice word of length \(T\).
For every deterministic machine \(M\) over the binary alphabet and every input \(x\), the blank-tape initial configuration of the guessing-phase machine of \(M\) on \(x\) equals the invariant configuration for the initial configuration of \(M\) on \(x\) with empty captured word.
Let \(L\) be a language over the binary alphabet and \(C, c\) natural numbers. If some nondeterministic machine decides \(L\) within time \(K\, (n + C(n+1)^c + 1)^r\) for some constants \(K\) and \(r\), then \(L\) belongs to \(\bigcup _{e} \mathrm{NTIME}(n^e + 1)\).
Let \(M\) be a deterministic machine over the binary alphabet, \(d\) a configuration, \(u\) a captured word and \(w\) a word of choice bits. Running the guessing-phase machine of \(M\) under \(w\) from the invariant configuration for \((d,u)\) yields the invariant configuration for the \(|w|\)-step run of \(M\) from \(d\), with the selection of \(w\) along the emission mask appended to the captured word.
Let \(M\) be a deterministic machine over the binary alphabet, \(d\) one of its configurations, \(u\) a captured word and \(b\) a choice bit. One nondeterministic step of the guessing-phase machine of \(M\) under \(b\) maps the invariant configuration for \((d, u)\) to the invariant configuration for the stepped configuration of \(M\), with \(b\) appended to the captured word exactly when \(M\) emits an output symbol at \(d\).
Let \(N\) be a nondeterministic machine and consider a loader configuration over a physical input \(y\) whose input head is at position \(i + 1\) with \(i \le |y|\). The symbol read from the physical input is the optional bit of \(y\) at index \(i\), which is blank exactly at the right boundary.
Let \(N\) be a nondeterministic machine, let \(y\) be a physical input with position index \(i \lt |y|\), and suppose the transition table of the pair-loader machine, applied in a loader state \(q\) to the bit of \(y\) at index \(i\), yields a right input move that optionally appends one bit to each buffer and enters a loader state \(q'\). Then one physical step from the loader configuration at position \(i+1\) with buffers \(x\) and \(u\) (heads at their right blanks) produces the loader configuration at position \(i+2\) in state \(q'\) with the appended buffers.
Let \(M\) be a deterministic machine, \(d\) a configuration and \(t\) a step count. The output length of \(d\) plus the number of \(1\) entries in the emission mask over \(t\) steps equals the output length of the configuration reached after \(t\) steps of \(M\) from \(d\).
Let \(N\) be a nondeterministic Turing machine over the binary alphabet and let \(x\) and \(u\) be binary strings. On the self-delimiting encoding of the pair \((x,u)\), the pair-loader machine of \(N\) halts from blank tapes within \(3(\ell + 1)\) steps, where \(\ell \) is the length of the encoded pair, and outputs the single bit recording whether the run of \(N\) on \(x\) under choice word \(u\) has halted with output the one-bit string \(1\).
Let \(N\) be a nondeterministic machine and suppose the physical input has the form \(y = p\, b\, b\, r\) for a bit \(b\). From the initial parser state at position \(|p|+1\) with input buffer \(x\), exactly two steps of the pair-loader machine append \(b\) to the input buffer and move to position \(|p|+3\), leaving the choice buffer empty and the output silent.
Let \(N\) be a nondeterministic machine and suppose the physical input has the form \(y = p\, e\) where \(e\) is the self-delimiting encoding of a pair \((x, u)\). From the initial parser state at position \(|p|+1\) with input buffer \(a\), exactly \(2|x| + 2\) steps of the pair-loader machine decode the doubled prefix and the separator, reaching the suffix-copy state at position \(|p| + 2|x| + 3\) with input buffer \(a\, x\) and the choice buffer still empty.
Let \(N\) be a nondeterministic machine and suppose the physical input has the form \(y = p\, 0\, 1\, u\). From the initial parser state at position \(|p|+1\) with input buffer \(x\), exactly two steps of the pair-loader machine enter the suffix-copy state at position \(|p|+3\) without modifying either buffer or emitting output.
For all natural numbers \(C\) and \(c\) there exist a deterministic scheduler machine \(M\) and a constant \(A\) such that for every \(n\), on the unary input of length \(n\) and with \(T = A(n+1)^{c+1}\): every choice word of length \(T\) drives the guessing-phase machine of \(M\) to the completed \(T\)-step configuration while capturing some word of length exactly \(C(n+1)^c\), and every word of that length is captured by some choice word of length \(T\).
Let \(N\) be a nondeterministic machine, let \(x\) and \(u\) be buffer words, and let \(j \le |u|\). From the loader configuration in the choice-rewind state with choice head at position \(j - 1\), exactly \(j + 1\) steps of the pair-loader machine reach the renamed embedding of the initial configuration of \(N\) on \(x\) prepared with choice word \(u\), preserving the physical input position.
Let \(N\) be a nondeterministic machine, let \(x\) and \(u\) be buffer words, and let \(j \le |x|\). From the loader configuration in the input-rewind state with input-buffer head at position \(j-1\) and choice head at the right end of \(u\), exactly \(j+1\) steps of the pair-loader machine reach the choice-rewind state with input-buffer head reset to zero and choice head at position \(|u| - 1\), with all tape contents preserved.
Let \(N\) be a nondeterministic Turing machine over the binary alphabet and let \(C, c\) be natural numbers. For every binary string \(y\), the pair-loader machine of \(N\), run on the word emitted by the polynomial split emitter for \(y\), halts within \(3(2|y|+3)\) steps, and its output is the indicator bit of membership of \(y\) in the choice-word verifier language of \(N\) with parameters \(C, c\).
Let \(N\) be a nondeterministic machine and let \(x\) and \(u\) be the buffer words held by a loader configuration at the right boundary of the physical input \(y\). Exactly \(1 + (|x|+1) + (|u|+1)\) further steps of the pair-loader machine (one dispatch and the two buffer rewinds) reach the renamed embedding of the initial configuration of \(N\) on \(x\) prepared with choice word \(u\); no bits of the choice word are consumed by these administrative steps.
Let \(N\) be a nondeterministic machine and let the physical input decompose as \(y = p\, u\). Starting in the suffix-copy state at position \(|p|+1\) with input buffer \(x\) and empty choice buffer, after \(|u| + (1 + (|x|+1) + (|u|+1)) + (|u|+1)\) steps the pair-loader machine has halted, and its output is the single bit recording whether the run of \(N\) on \(x\) under choice word \(u\) has halted with output the one-bit string \(1\).
Let \(N\) be a nondeterministic machine and consider a loader configuration whose two buffer heads stand at the right blanks of the buffer words \(x\) and \(u\). Applying an action that optionally writes one bit at each of the two buffer heads (advancing a head exactly when it writes) produces the loader configuration whose buffers are \(x\) and \(u\) with the respective optional bits appended, with the physical input head moved by the action’s input move and all other tapes unchanged.
For a deterministic machine \(M\), a configuration \(d\) and a step count \(t\), the emission mask is the length-\(t\) bit list whose \(i\)-th entry records whether the machine emits an output symbol on the \((i+1)\)-st step of its run from \(d\).
Given a configuration \(d\) of a deterministic machine \(M\) and a captured word \(u\), the corresponding configuration of the guessing-phase machine of \(M\) carries the state, input head and work tapes of \(d\) unchanged, holds \(u\) on the extra capture tape with its head at the right blank, and has empty physical output.
For a deterministic machine \(M\) over the binary alphabet, a nondeterministic machine with one extra tape is defined: it runs \(M\) step for step, and whenever \(M\) emits an output symbol it instead appends the current nondeterministic choice bit to the extra capture tape; when \(M\) halts, the phase stays in a live return state and makes no further changes.
For a nondeterministic machine \(N\), a loader state index, two buffer words \(x\) and \(u\), a physical input position and two buffer head positions, the corresponding configuration of the pair-loader machine of \(N\) holds \(x\) and \(u\) on the two buffer tapes, with all source work tapes and the capture tape blank and the physical output empty.
For a nondeterministic Turing machine \(N\) over the binary alphabet, a deterministic machine is defined that parses a self-delimiting encoded pair on its input tape: it decodes the doubled-bit prefix into an input buffer, copies the suffix after the separator into a choice buffer, rewinds both buffers, and then runs the simulation core of \(N\) with its transition table embedded unchanged. An input that is not an aligned pair (in particular the empty word) is rejected with output bit \(0\).
Given a mask and a word of physical choice bits, processed position by position, the selected word keeps the choice bit at every position where the mask is \(1\) and drops it where the mask is \(0\); positions beyond the shorter of the two lists contribute nothing.
For natural numbers \(C, c\) and a binary string \(y\), the emitted word is the self-delimiting encoding of the pair formed by the first \(i\) bits of \(y\) and the remaining bits of \(y\), when the split search solves \(i + C\, (i+1)^c = |y|\); it is the empty word when the search fails.
For a binary machine \(M\) with \(k\) work tapes, a word \(x\), and a count \(j\), the copied payload at coordinate \(z\) is blank on every work lane; on the input lane it is the virtual input payload of \(x\) at \(z\) when \(z = -1\) (the copied left boundary) or \(0 \le z \lt j\), and blank otherwise.
For every binary machine \(M\) and word \(x\), the copied payloads of count \(|x| + 1\) coincide with the head-centered payload vectors of the initial configuration of \(M\) on \(x\): the input lane carries the complete virtual input including both boundary tags, and every work lane is blank.
Let \(x\) be a binary word embedded into the simulation alphabet and let a configuration on the embedded input have its input head at position \(j + 1\) with \(j \le |x|\). Then the payload whose content is the clock-bit decoding of the input read and whose tag is the right-boundary tag \(2\) when the read is blank and the ordinary tag \(0\) otherwise is exactly the virtual input payload of \(x\) at position \(j\).
Let \(x\) be a binary word embedded letterwise into the simulation alphabet, and let a configuration on this embedded input have its input head at position \(j + 1\) with \(j \le |x|\). Then the input read is the bit symbol of \(x_j\) when \(j \lt |x|\), and blank when \(j = |x|\); this includes the empty input.
For a binary machine \(M\), given contents of the data tapes, a coordinate \(z\), and a payload \(p\), the copy write updates only the input lane, placing at \(z\) a data cell with current payload \(p\) and blank cached neighbor; the work lanes are unchanged.
If the current payloads of the data tapes are the copied payloads of count \(j\), then after the copy write placing the virtual input payload of \(x\) at position \(j\) at coordinate \(j\), the current payloads are the copied payloads of count \(j + 1\).
For \(j \in \mathbb {N}\), the partial guide after copying \(j\) input cells is the tape whose cell \(-1\) carries an interior mark (the virtual left boundary), whose cells \(z\) with \(0 \le z \lt j\) carry the origin mark at \(z = 0\) and interior marks otherwise, and whose remaining cells are blank.
Suppose the four-state binary-counter machine is in its carry state with the concatenation of binary words \(p\) and \(b\) on the nonnegative work-tape cells, work head at cell \(|p|\), and empty output, and let \(r\) be the carry length of \(b\) (its number of leading \(1\) bits). Then after exactly \(r + 1\) steps the machine is in the rewind state with the work tape holding \(p\) followed by the incremented word of \(b\), the work head at cell \(|p| + r - 1\), and the input head and output unchanged.
Suppose the four-state binary-counter machine is in its carry state with the concatenation of binary words \(p\) and \(b\) on the nonnegative work-tape cells, work head at cell \(|p|\), and empty output. If \(b\) begins with \(1\), one step clears that bit to \(0\) and moves the work head right, staying in the carry state; otherwise one step writes \(1\) at cell \(|p|\) (replacing the first bit of \(b\), or extending the word) and moves the work head left, entering the rewind state. The input head does not move.
Let \(x\) be a binary input word and \(i \le |x|\), and let \(w(i)\) denote the number of \(1\) bits in the binary representation of \(i\). Then there is a step count \(t\) with \(t + 2\, w(i) \le 4i\) such that, after \(t\) steps from its initial configuration on \(x\), the four-state binary-counter machine is in the count state having consumed exactly the first \(i\) input symbols, with the work tape holding the binary representation of \(i\) (least-significant bit first), the work head at cell \(0\), and empty output.
Suppose the four-state binary-counter machine is in its emit state with a binary word \(b\) on the nonnegative work-tape cells, work head at cell \(0\), and empty output. Then after exactly \(|b| + 1\) steps the machine has halted and its output equals \(b\).
Suppose the four-state binary-counter machine is in its emit state with a binary word \(b\) on the nonnegative work-tape cells, work head at cell \(0\), and empty output. Then for every \(i \le |b|\), after exactly \(i\) steps it is still in the emit state with the work head at cell \(i\), the tape contents and input-head position unchanged, and the output equal to the first \(i\) bits of \(b\).
Let \(x\) be an input word and \(i \lt |x|\). Starting the four-state binary-counter machine in the count state with the input head scanning the \(i\)-th input symbol, a binary word \(b\) stored on the work tape with the work head at cell \(0\), and empty output: after exactly \(2r + 2\) steps, where \(r\) is the carry length of \(b\), the machine is again in the count state with the input head advanced by one position, the work tape holding the incremented word of \(b\), the work head back at cell \(0\), and the output still empty.
Suppose the four-state binary-counter machine is in its rewind state with a binary word \(b\) on the nonnegative work-tape cells, empty output, and work head at cell \(j - 1\) for some natural number \(j \le |b|\) (so possibly at the blank cell \(-1\)). Then after exactly \(j + 1\) steps the machine is in the count state with the work head at cell \(0\), and the stored word, input-head position, and output all unchanged.
Let \(x\) be an input word and \(i \lt |x|\). Taking one step of the four-state binary-counter machine from the count state — with the input head scanning the \(i\)-th input symbol, the work head at cell \(0\) over a stored binary word \(b\), and empty output — produces the same configuration as taking one step from the carry state with the input head advanced by one position and the same work tape and output.
Given an input-head displacement \(d \in \{ -1, 0, +1\} \) and the symbol currently scanned on the work tape, the action of one carry step of a one-work-tape binary machine: if the scanned cell holds \(1\), it is overwritten with \(0\), the work head moves right, and the machine continues in its carry state; otherwise (a \(0\) or a blank) the cell is overwritten with \(1\), the work head moves left, and the machine enters its rewind state. In both cases the input head moves by \(d\) and nothing is written to the output.
The configuration of the one-work-tape four-state counter machine on an input word \(x\) specified by: a state \(q\) among the four counter states, an input-head position \(p\), an integer work-head position \(z\), a finite binary word \(b\) laid out on the nonnegative cells of the work tape (blank elsewhere), and an output word.
The successor operation on binary words written least-significant bit first: the empty word maps to the single-bit word \(1\); a word beginning with \(0\) has that bit set to \(1\) and the rest unchanged; a word beginning with \(1\) has that bit cleared to \(0\) and the increment carries recursively into the remaining word.
For every natural number \(n\), applying the word increment to the little-endian binary representation of \(n\) (with no redundant leading zeros; \(0\) is the empty word) yields exactly the binary representation of \(n + 1\), including the case where the word grows by one bit on overflow.
Let \(b\) be a binary word written least-significant bit first, let \(r\) be its carry length (its number of leading \(1\) bits), and let \(w(b)\) denote the number of \(1\) bits in a word. Then the incremented word \(b'\) satisfies \(w(b') + r = w(b) + 1\): each cleared bit lowers the count of ones by one, and the single final write raises it by one.
Consider machines with \(R\) work tapes over the bit alphabet whose control states are the compiled control states for \(R\) registers and labels \(\Lambda \). Given a work tape index \(t\), an input-head move \(i \in \{ -1, 0, 1\} \), a write instruction (write nothing, write a blank, or write a bit), a head move \(m \in \{ -1, 0, 1\} \), an optional output bit \(o\) and an optional next control state \(q\), this is the single-step action that moves the input head by \(i\), performs the write instruction on tape \(t\) and moves its head by \(m\), leaves every other work tape untouched (no write, no movement), emits \(o\) if present, and passes to \(q\) (halting if \(q\) is absent).
Let \(x\) be an input bit string and consider a configuration on \(x\), with \(R\) work tapes and the compiled control states, having an arbitrary state, input-head position \(p\), work tapes \((T_i)\), work heads \((h_i)\) and output \(w\). Applying to it the action on the single work tape \(t\) with input-head move \(i\), write instruction \(\omega \), head move \(m\), optional output bit \(o\) and optional next state \(q\) yields the configuration with state \(q\), input-head position the clamped move of \(p\) by \(i\), tape \(t\) replaced by \(T_t\) with cell \(h_t\) overwritten as prescribed by \(\omega \) (and unchanged if \(\omega \) writes nothing), head \(t\) moved to \(h_t + m\), all other tapes and heads unchanged, and output \(w\) followed by \(o\) if \(o\) is present.
Let \(x\) be an input bit string and consider a configuration on \(x\), with \(R\) work tapes and the compiled control states, having an arbitrary state, input-head position \(p\), work tapes \((T_i)\), work heads \((h_i)\) and output \(w\). Applying to it the action that touches no work tape, with input-head move \(i\), optional output bit \(o\) and optional next state \(q\), yields the configuration with state \(q\), input-head position the clamped move of \(p\) by \(i\), the same work tapes \((T_i)\) and heads \((h_i)\), and output \(w\) followed by \(o\) if \(o\) is present.
Let \(\Lambda \) be a finite set of labels, \(P\) a counter program over \(R\) registers with labels in \(\Lambda \), \(\ell _0 \in \Lambda \), \(x\) an input bit string, \(r\) a register, \(\ell \) a label, \(p\) an input-head position, \((T_i)\) work tapes, \((h_i)\) work heads, \(w\) an output string, and \(m \in \mathbb {N}\) such that tape \(r\) is the unary tape \(T_r = 1^m\) (cells \(0, \dots , m-1\) hold \(1\), all others blank). Then for every \(d \le m\), running the machine of \(P\) started at \(\ell _0\) for \(d + 1\) steps from the configuration in state \(\mathsf{print}_{\rightarrow }(r, \ell )\) with input head \(p\), tapes \((T_i)\), heads \((h_i)\) except that head \(r\) is on cell \(m - d\), and output \(w\), reaches the configuration in state \(\mathsf{main}(\ell )\) with input head \(p\), tapes \((T_i)\), heads \((h_i)\) except that head \(r\) is on cell \(m\), and output \(w\).
Consider machines with \(R\) work tapes over the bit alphabet whose control states are the compiled control states for \(R\) registers and labels \(\Lambda \). Given an input-head move \(i \in \{ -1, 0, 1\} \), an optional output bit \(o\) and an optional next control state \(q\), this is the single-step action that moves the input head by \(i\), neither writes on nor moves any work tape, emits \(o\) if present, and passes to \(q\) (halting if \(q\) is absent).
For a natural number \(v\) write \(1^v\) for the unary tape whose cells \(0, \dots , v-1\) hold \(1\) and whose other cells (indexed by \(\mathbb {Z}\)) are blank. Let \(x\) be an input bit string of length \(n\) and \(s\) an abstract state with registers \(\rho \), read position \(p\) and output \(w\). Its encoding is the configuration, on input \(x\), of a machine with \(R\) work tapes and the compiled control states, whose state is \(\mathsf{main}(\ell )\) if \(s\) has label \(\ell \) and is halted if \(s\) is halted, whose input head is at position \(\min (p + 1, n + 1)\), whose work tape \(r\) holds \(1^{\rho (r)}\) with its head on cell \(\rho (r)\), and whose output tape holds \(w\).
Let \(x\) be an input bit string of length \(n\) and \(s\) an abstract state of a counter program over \(R\) registers whose read position \(p\) satisfies \(p \le n\). Then in the machine configuration encoding \(s\) on input \(x\), the input head is at position \(p + 1\).
Let \(x\) be an input bit string of length \(n\) and \(s\) an abstract state of a counter program over \(R\) registers whose read position \(p\) satisfies \(p \le n\). Then the symbol under the input head in the machine configuration encoding \(s\) on input \(x\) is the bit \(x_p\) (indexing from \(0\)) when \(p \lt n\), and no symbol (a blank) when \(p = n\).
Let \(P\) be a counter program over \(R\) registers with a finite label set \(\Lambda \) and start label \(\ell _0\), let \(x\) be a binary input and \(t \in \mathbb {N}\). Suppose that \(t\) steps of \(P\) on \(x\), started from the initial state (label \(\ell _0\), all registers \(0\), nothing read, empty output), reach a halted state, and let \(y\) be its output. Then the multi-tape Turing machine compiled from \(P\) with start label \(\ell _0\), which keeps each register on its own work tape in unary, computes \(y\) on input \(x\) in time \(t(2t+3)\): after \(t(2t+3)\) steps from its initial configuration on \(x\) it is halted with output \(y\).
For any register values \(\rho \) and any bit string \(w = w_1 \cdots w_n\), concatenating the strings printed under \(\rho \) by the micro-operations \(\mathsf{out}(w_1), \dots , \mathsf{out}(w_n)\), where \(\mathsf{out}(b)\) prints the single bit \(b\), gives back \(w\).
Let \(\rho : \{ 0,\dots ,R-1\} \to \mathbb {N}\) be register values and \((L, c)\) a linear expression with \(L = (r_1, \dots , r_k)\). Concatenating the strings printed under \(\rho \) by the micro-operations \(\mathsf{pr}(r_1), \dots , \mathsf{pr}(r_k)\) followed by \(c\) copies of \(\mathsf{out}(1)\), where \(\mathsf{pr}(r)\) prints \(1^{\rho (r)}\) and \(\mathsf{out}(1)\) prints \(1\), gives exactly the unary string \(1^{v}\) with \(v = \rho (r_1) + \dots + \rho (r_k) + c\).
Let \(P\) be a counter program over \(R\) natural-number registers with label set \(\Lambda \), run on a binary input \(x\). Given a label \(\ell \in \Lambda \), a target \(\ell ' \in \Lambda \cup \{ \bot \} \) (where \(\bot \) means halted), register vectors \(\rho , \rho ' : \{ 0,\dots ,R-1\} \to \mathbb {N}\), input positions \(p, p' \in \mathbb {N}\), a bit string \(e\) and a bound \(b \in \mathbb {N}\), this is the predicate that for every output \(o\) printed so far there is some \(t \le b\) such that \(t\) steps of \(P\) on \(x\) lead from the state with label \(\ell \), registers \(\rho \), \(p\) input symbols read and output \(o\) to the state with label \(\ell '\), registers \(\rho '\), \(p'\) input symbols read and output \(o\, e\).
Let \(P\) be a counter program on a binary input \(x\). Suppose that from label \(\ell \) with registers \(\rho \) and \(p\) input symbols read, \(P\) reaches \(\ell '\) (a label or the halted state) with registers \(\rho '\) and \(p'\) symbols read within \(b\) steps while appending \(e\) to its output, and that \(\rho = \sigma \), \(\rho ' = \sigma '\), \(p = q\), \(p' = q'\), \(e = e'\) and \(b \le b'\). Then from label \(\ell \) with registers \(\sigma \) and \(q\) symbols read, \(P\) reaches \(\ell '\) with registers \(\sigma '\) and \(q'\) symbols read within \(b'\) steps while appending \(e'\).
Let \(P\) be a counter program on a binary input \(x\) and let \(b \le b'\). If from label \(\ell \) with registers \(\rho \) and \(p\) input symbols read, \(P\) reaches \(\ell '\) (a label or the halted state) with registers \(\rho '\) and \(p'\) symbols read within \(b\) steps while appending \(e\) to its output, then it does so within \(b'\) steps as well.
Let \(P\) be a counter program on a binary input \(x\), let \(\ell , \ell '\) be labels, \(\ell ''\) a label or the halted state, \(\rho , \rho ', \rho ''\) register vectors and \(p, p', p''\) input positions. Suppose that from label \(\ell \) with registers \(\rho \) and \(p\) symbols read, \(P\) reaches label \(\ell '\) with registers \(\rho '\) and \(p'\) symbols read within \(b\) steps while appending \(e\) to whatever output was printed before, and that from \((\ell ', \rho ', p')\) it reaches \((\ell '', \rho '', p'')\) within \(b'\) steps while appending \(e'\). Then from \((\ell , \rho , p)\) it reaches \((\ell '', \rho '', p'')\) within \(b + b'\) steps while appending the concatenation \(e\, e'\).
Let \(P\) be a counter program on a binary input \(x\), \(r\) a register and \(\ell \) a label whose instruction is \(\mathsf{dec}(r, \ell ')\). Then for all registers \(\rho \) and input positions \(p\), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell '\) within one step, printing nothing, with \(p\) symbols read and registers \(\rho \) except that register \(r\) now holds \(\max (\rho (r) - 1, 0)\).
Let \(P\) be a counter program on a binary input \(x\) and \(\ell \) a label whose instruction is \(\mathsf{goto}\ \ell '\). Then for all registers \(\rho \) and input positions \(p\), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell '\) with registers \(\rho \) and \(p\) symbols read within one step, without printing anything.
Let \(P\) be a counter program on a binary input \(x\) and \(\ell \) a label whose instruction is \(\mathsf{halt}\). Then for all registers \(\rho \) and input positions \(p\), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches the halted state, with registers still \(\rho \) and still \(p\) symbols read, within one step and without printing anything.
Let \(P\) be a counter program on a binary input \(x\), \(r\) a register and \(\ell \) a label whose instruction is \(\mathsf{inc}(r, \ell ')\). Then for all registers \(\rho \) and input positions \(p\), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell '\) within one step, printing nothing, with \(p\) symbols read and registers \(\rho \) except that register \(r\) now holds \(\rho (r) + 1\).
Let \(P\) be a counter program on a binary input \(x\), \(r\) a register and \(\ell \) a label whose instruction is \(\mathsf{jz}(r, \ell _0, \ell _1)\), “go to \(\ell _0\) if register \(r\) is zero, else to \(\ell _1\)”. If the registers \(\rho \) satisfy \(\rho (r) \ne 0\), then for every input position \(p\), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell _1\) with registers \(\rho \) and \(p\) symbols read within one step, printing nothing.
Let \(P\) be a counter program on a binary input \(x\), \(r\) a register and \(\ell \) a label whose instruction is \(\mathsf{jz}(r, \ell _0, \ell _1)\), “go to \(\ell _0\) if register \(r\) is zero, else to \(\ell _1\)”. If the registers \(\rho \) satisfy \(\rho (r) = 0\), then for every input position \(p\), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell _0\) with registers \(\rho \) and \(p\) symbols read within one step, printing nothing.
Let \(P\) be a counter program on a binary input \(x\), \(r\) a register, and \(h, d, c, z\) labels such that the instruction at \(h\) is \(\mathsf{jz}(r, z, d)\) (go to \(z\) if register \(r\) is zero, else to \(d\)) and the instruction at \(d\) is \(\mathsf{dec}(r, c)\). Let \(K, b \in \mathbb {N}\), and for each \(j \in \mathbb {N}\) let \(\sigma _j\) be a register vector, \(p_j\) an input position and \(e_j\) a bit string, with \(\sigma _j(r) = K - j\) for all \(j \le K\). Suppose that for every \(j \lt K\), from label \(c\) with registers \(\sigma _j\) modified so that register \(r\) holds \(K - j - 1\), and \(p_j\) symbols read, \(P\) reaches label \(h\) with registers \(\sigma _{j+1}\) and \(p_{j+1}\) symbols read within \(b\) steps while appending \(e_j\) to its output. Then from label \(h\) with registers \(\sigma _0\) and \(p_0\) symbols read, \(P\) reaches label \(z\) with registers \(\sigma _K\) and \(p_K\) symbols read within \(K(b+2)+1\) steps while appending the concatenation \(e_0 e_1 \cdots e_{K-1}\).
Let \(P\) be a counter program on a binary input \(x\), let \(\ell \) be a label, \(\ell '\) a label or the halted state, \(\rho , \rho '\) register vectors, \(p, p'\) input positions and \(e\) a bit string. If for every output \(o\) one step of \(P\) on \(x\) maps the state with label \(\ell \), registers \(\rho \), \(p\) symbols read and output \(o\) to the state with label \(\ell '\), registers \(\rho '\), \(p'\) symbols read and output \(o\, e\), then from \((\ell , \rho , p)\) the program reaches \((\ell ', \rho ', p')\) within one step while appending \(e\) to its output.
Let \(P\) be a counter program on a binary input \(x\) and \(\ell \) a label whose instruction is \(\mathsf{out}(b, \ell ')\), “print the bit \(b\), then go to \(\ell '\)”. Then for all registers \(\rho \) and input positions \(p\), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell '\) with registers \(\rho \) and \(p\) symbols read within one step, appending the single bit \(b\) to its output.
Let \(P\) be a counter program on a binary input \(x\), \(r\) a register and \(\ell \) a label whose instruction is \(\mathsf{pr}(r, \ell ')\), “print the value of register \(r\) in unary, then go to \(\ell '\)”. Then for all registers \(\rho \) and input positions \(p\), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell '\) with registers \(\rho \) and \(p\) symbols read within one step, appending the string \(1^{\rho (r)}\) of \(\rho (r)\) ones to its output.
Let \(P\) be a counter program on a binary input \(x\) and \(\ell \) a label whose instruction is \(\mathsf{rd}(\ell _e, \ell _0, \ell _1)\), “read the next input symbol and go to \(\ell _e\), \(\ell _0\) or \(\ell _1\) according as the input has ended, the symbol is \(0\), or it is \(1\)”. If \(x\) has no symbol at (zero-based) position \(p\), i.e. \(p \ge |x|\), then for all registers \(\rho \), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell _e\) with registers \(\rho \) and still \(p\) symbols read within one step, printing nothing.
Let \(P\) be a counter program on a binary input \(x\) and \(\ell \) a label whose instruction is \(\mathsf{rd}(\ell _e, \ell _0, \ell _1)\), “read the next input symbol and go to \(\ell _e\), \(\ell _0\) or \(\ell _1\) according as the input has ended, the symbol is \(0\), or it is \(1\)”. If the symbol of \(x\) at (zero-based) position \(p\) is \(0\), then for all registers \(\rho \), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell _0\) with registers \(\rho \) and \(p + 1\) symbols read within one step, printing nothing.
Let \(P\) be a counter program on a binary input \(x\) and \(\ell \) a label whose instruction is \(\mathsf{rd}(\ell _e, \ell _0, \ell _1)\), “read the next input symbol and go to \(\ell _e\), \(\ell _0\) or \(\ell _1\) according as the input has ended, the symbol is \(0\), or it is \(1\)”. If the symbol of \(x\) at (zero-based) position \(p\) is \(1\), then for all registers \(\rho \), from label \(\ell \) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\ell _1\) with registers \(\rho \) and \(p + 1\) symbols read within one step, printing nothing.
Let \(P\) be a counter program on a binary input \(x\), let \(\tau = (m_0, \dots , m_{n-1})\) be a list of micro-operations, each either “print a bit \(b\)” or “print register \(r\) in unary”, and let \(\lambda : \mathbb {N} \to \Lambda \) and \(\nu \in \Lambda \) be labels. Suppose that for each \(i \lt n\) the instruction of \(P\) at \(\lambda (i)\) performs \(m_i\) and then goes to \(\lambda (i+1)\), and the instruction at \(\lambda (n)\) is \(\mathsf{goto}\ \nu \). Then for all registers \(\rho \) and input positions \(p\), from label \(\lambda (0)\) with registers \(\rho \) and \(p\) symbols read the program reaches label \(\nu \) with registers \(\rho \) and \(p\) symbols read within \(n + 1\) steps, appending \(w_0 w_1 \cdots w_{n-1}\) to its output, where \(w_i\) is \(b\) if \(m_i\) prints the bit \(b\) and \(w_i = 1^{\rho (r)}\) if \(m_i\) prints register \(r\).
Let \(\rho : \{ 0, \dots , R-1\} \to \mathbb {N}\), let \(t\) be an index and \(m \in \mathbb {N}\), and let \(\rho '\) agree with \(\rho \) except that \(\rho '(t) = m\). Then replacing the \(t\)-th entry of the integer vector \((\rho (r))_r\) by \(m\) gives the integer vector \((\rho '(r))_r\).
Fix a number \(R\) of registers, indexed by \(\{ 0, \dots , R-1\} \), and a set \(\Lambda \) of labels. An instruction of a counter program over \(R\) registers with labels in \(\Lambda \) is one of the following, where \(r\) ranges over registers, \(b\) over bits and \(\ell ', \ell _0, \ell _+, \ell _\bot , \ell _1\) over labels: \(\mathsf{halt}\) (stop); \(\mathsf{goto}(\ell ')\) (jump to \(\ell '\)); \(\mathsf{out}(b, \ell ')\) (print the bit \(b\), then go to \(\ell '\)); \(\mathsf{inc}(r, \ell ')\) (increment register \(r\)); \(\mathsf{dec}(r, \ell ')\) (decrement register \(r\), with \(0\) staying \(0\)); \(\mathsf{jz}(r, \ell _0, \ell _+)\) (go to \(\ell _0\) if register \(r\) is \(0\), and to \(\ell _+\) otherwise); \(\mathsf{pr}(r, \ell ')\) (print \(1^v\), where \(v\) is the value of register \(r\)); and \(\mathsf{rd}(\ell _\bot , \ell _0, \ell _1)\) (read the next input symbol: go to \(\ell _\bot \) if the input is exhausted, and otherwise consume the symbol and go to \(\ell _0\) or \(\ell _1\) according as it is \(0\) or \(1\)).
For \(R \in \mathbb {N}\), a linear expression over \(R\) registers is a pair \((L, c)\) consisting of a finite list \(L\) of register indices in \(\{ 0, \dots , R-1\} \), repetitions allowed, and a constant \(c \in \mathbb {N}\); it stands for the sum of the listed registers plus \(c\).
For a linear expression \((L, c)\) over \(R\) registers with \(L = (r_1, \dots , r_k)\), this is the list of template micro-operations \(\mathsf{pr}(r_1), \dots , \mathsf{pr}(r_k)\) (print each listed register in unary) followed by \(c\) copies of \(\mathsf{out}(1)\) (print the bit \(1\)).
For \(R \in \mathbb {N}\), a template micro-operation over \(R\) registers is either \(\mathsf{out}(b)\) for a bit \(b\), meaning “print the bit \(b\)”, or \(\mathsf{pr}(r)\) for a register index \(r \lt R\), meaning “print the value of register \(r\) in unary”.
Given register values \(\rho : \{ 0,\dots ,R-1\} \to \mathbb {N}\), the string printed by a template micro-operation is defined by: \(\mathsf{out}(b)\) prints the one-symbol string \(b\), and \(\mathsf{pr}(r)\) prints \(1^{\rho (r)}\), the string of \(\rho (r)\) ones.
Let \(P\) be a counter program over \(R\) natural-number registers with a finite set of labels \(\Lambda \), i.e. an assignment to each label of one instruction (halt, jump, print a bit, increment a register, decrement a register saturating at \(0\), branch on a register being zero, print the value \(v\) of a register as \(1^v\), or read and consume the next input bit, branching on end-of-input, \(0\) or \(1\)). Fix a start label \(\ell _0 \in \Lambda \), a function \(f : \{ 0,1\} ^* \to \{ 0,1\} ^*\) and constants \(C, c \in \mathbb {N}\), and suppose that for every input \(x \in \{ 0,1\} ^*\) there is some \(t \le C(|x|+1)^c\) such that running \(P\) on \(x\) for \(t\) steps from the initial state (label \(\ell _0\), all registers \(0\), no input read, empty output) reaches a halted state whose output is \(f(x)\). Then \(f\) is polynomial-time computable: some finite binary-alphabet Turing machine computes \(f\) within \(C'(n+1)^{c'}\) steps on every input of length \(n\), for some constants \(C', c' \in \mathbb {N}\).
Let \(P\) be a counter program over \(R\) natural-number registers with a finite set of labels \(\Lambda \) (each label carrying one instruction: halt, jump, print a bit, increment, saturating decrement, zero-test branch, print a register’s value \(v\) as \(1^v\), or read and consume the next input bit), let \(\ell _0 \in \Lambda \), let \(f : \{ 0,1\} ^* \to \{ 0,1\} ^*\) and let \(C, c \in \mathbb {N}\). Suppose that for every input \(x \in \{ 0,1\} ^*\) there are register values \(\rho : \{ 0,\dots ,R-1\} \to \mathbb {N}\) and a position \(p \in \mathbb {N}\) such that, for every previously printed output \(o\), running \(P\) on \(x\) from label \(\ell _0\) with all registers \(0\), no input read and output \(o\) reaches, within at most \(C(|x|+1)^c\) steps, the halted state with registers \(\rho \), \(p\) input bits read and output \(o\) followed by \(f(x)\). Then \(f\) is polynomial-time computable: some finite binary-alphabet Turing machine computes \(f\) within \(C'(n+1)^{c'}\) steps on every input of length \(n\), for some constants \(C', c' \in \mathbb {N}\).
Let \(\Lambda \) be a finite set of labels, \(P\) a counter program over \(R\) registers with labels in \(\Lambda \), \(\ell _0 \in \Lambda \), \(x\) an input bit string, \(r\) a register, \(\ell \) a label, \(p\) an input-head position, \((T_i)\) work tapes and \((h_i)\) work heads, and \(m \in \mathbb {N}\) such that tape \(r\) is the unary tape \(T_r = 1^m\) (cells \(0, \dots , m-1\) hold \(1\), all others blank). Then for every \(j \le m\) and every output string \(w\), running the machine of \(P\) started at \(\ell _0\) for \(j + 1\) steps from the configuration in state \(\mathsf{print}_{\leftarrow }(r, \ell )\) with input head \(p\), tapes \((T_i)\), heads \((h_i)\) except that head \(r\) is on cell \(j - 1\), and output \(w\), reaches the configuration in state \(\mathsf{print}_{\rightarrow }(r, \ell )\) with input head \(p\), tapes \((T_i)\), heads \((h_i)\) except that head \(r\) is on cell \(0\), and output \(w\, 1^j\).
Given a counter program \(P\) over \(R\) registers with labels in \(\Lambda \), an input bit string \(x\), an abstract state \(s\) and \(t \in \mathbb {N}\), the \(t\)-step run \(\mathrm{run}_{P,x}(s, t)\) is the state obtained by applying the one-step map of \(P\) on \(x\) to \(s\) exactly \(t\) times.
Let \(P\) be a counter program over \(R\) registers with labels in \(\Lambda \), \(x\) an input bit string and \(s\) an abstract state. For all \(a, b \in \mathbb {N}\), running \(a + b\) steps from \(s\) is the same as running \(a\) steps and then \(b\) further steps: \(\mathrm{run}_{P,x}(s, a + b) = \mathrm{run}_{P,x}(\mathrm{run}_{P,x}(s, a), b)\).
Let \(P\) be a counter program on a binary input \(x\), \(\ell _0\) a label and \(t \in \mathbb {N}\). Starting from the initial state at \(\ell _0\) (all registers \(0\), nothing read, empty output), the output after \(t\) steps of \(P\) on \(x\) has length at most \(t(t+1)\).
Let \(P\) be a counter program over \(R\) registers with labels in \(\Lambda \), \(x\) an input bit string, and \(s\) an abstract state whose label component marks it as halted. Then for every \(t \in \mathbb {N}\), running \(P\) on \(x\) from \(s\) for \(t\) steps returns \(s\) unchanged.
Let \(P\) be a counter program on a binary input \(x\), \(s\) an abstract state of \(P\) with output \(o\), and \(t \in \mathbb {N}\). Then there is a bit string \(e\) such that the output after \(t\) steps of \(P\) on \(x\) from \(s\) is \(o\, e\).
Let \(P\) be a counter program on a binary input \(x\), let \(s\) be an abstract state of \(P\) in which \(\mathrm{pos}(s)\) input symbols have been read, and let \(t \in \mathbb {N}\). Then after \(t\) steps of \(P\) on \(x\) from \(s\) the number of input symbols read is at most \(\mathrm{pos}(s) + t\).
Let \(P\) be a counter program over \(R\) registers with labels in \(\Lambda \), \(x\) an input bit string, \(s\) an abstract state with registers \(\rho \), \(r\) a register and \(t \in \mathbb {N}\). If \(\rho '\) denotes the registers of the state reached from \(s\) after \(t\) steps of \(P\) on \(x\), then \(\rho '(r) \le \rho (r) + t\).
Let \(P\) be a counter program over \(R\) registers with labels in \(\Lambda \), \(x\) an input bit string, \(s\) an abstract state and \(t \in \mathbb {N}\). Running \(t + 1\) steps from \(s\) equals running \(t\) steps from the state obtained by one step of \(P\) on \(x\) from \(s\): \(\mathrm{run}_{P,x}(s, t + 1) = \mathrm{run}_{P,x}(\mathrm{step}_{P,x}(s), t)\).
Let \(P\) be a counter program over \(R\) registers with a finite label set \(\Lambda \), let \(\ell _0 \in \Lambda \), let \(x\) be a binary input, and let \(M\) be the multi-tape Turing machine compiled from \(P\) with start label \(\ell _0\), which keeps each register on its own work tape in unary. For an abstract state \(s\) (label or halted, registers \(\rho _s\), number \(\mathrm{pos}(s)\) of input symbols read, output so far) let \(\mathrm{enc}(s)\) be the configuration of \(M\) in the control state for the label of \(s\) (or halted), with the input head on the next unread symbol, work tape \(r\) holding \(1^{\rho _s(r)}\) with its head on cell \(\rho _s(r)\), and output equal to that of \(s\). Let \(B \in \mathbb {N}\). Then for every \(t \in \mathbb {N}\) and every abstract state \(s\) with \(\mathrm{pos}(s) \le |x|\) and \(\rho _s(r) + t \le B\) for every register \(r\), there is \(t' \le t(2B+3)\) such that running \(M\) for \(t'\) steps from \(\mathrm{enc}(s)\) yields \(\mathrm{enc}(s_t)\), where \(s_t\) is the state reached by \(t\) steps of \(P\) on \(x\) from \(s\).
Let \(\Lambda \) be a finite set of labels, \(P\) a counter program over \(R\) registers with labels in \(\Lambda \), \(\ell _0 \in \Lambda \), \(x\) an input bit string, and \(B \in \mathbb {N}\). Let \(s\) be an abstract state that is running at some label \(\ell \in \Lambda \), whose read position is at most \(|x|\), and all of whose registers are at most \(B\). Then there is \(t \le 2B + 3\) such that running the Turing machine of \(P\) started at \(\ell _0\) for \(t\) steps from the configuration encoding \(s\) on input \(x\) yields the configuration encoding the state reached from \(s\) by one step of \(P\) on \(x\). Here a state with registers \(\rho \), read position \(p\) and output \(w\) is encoded by the configuration in control state \(\mathsf{main}(\ell )\) (or halted if the state is halted) whose input head is at position \(\min (p + 1, |x| + 1)\), whose work tape \(r\) holds \(1\) exactly on cells \(0, \dots , \rho (r) - 1\) with its head on cell \(\rho (r)\), and whose output is \(w\).
For \(R\) registers and a set \(\Lambda \) of labels, an abstract state of a counter program consists of four components: the current label, which is either an element of \(\Lambda \) or a special value marking that the program has halted; a register assignment \(\rho : \{ 0, \dots , R-1\} \to \mathbb {N}\); the read position \(p \in \mathbb {N}\), the number of input symbols consumed so far; and the output printed so far, a finite bit string.
Let \(P : \Lambda \to \mathrm{Instr}\) be a counter program over \(R\) registers (an instruction at every label) and \(x\) an input bit string. One step of \(P\) on \(x\) leaves a halted state unchanged. A state with label \(\ell \), registers \(\rho \), read position \(p\) and output \(w\) is updated according to \(P(\ell )\), every component not mentioned staying the same: \(\mathsf{halt}\) makes the state halted; \(\mathsf{goto}(\ell ')\) moves to label \(\ell '\); \(\mathsf{out}(b, \ell ')\) moves to \(\ell '\) and replaces the output by \(w b\); \(\mathsf{inc}(r, \ell ')\) and \(\mathsf{dec}(r, \ell ')\) move to \(\ell '\) and replace \(\rho (r)\) by \(\rho (r) + 1\), respectively by the truncated difference \(\max (\rho (r) - 1, 0)\); \(\mathsf{jz}(r, \ell _0, \ell _+)\) moves to \(\ell _0\) if \(\rho (r) = 0\) and to \(\ell _+\) otherwise; \(\mathsf{pr}(r, \ell ')\) moves to \(\ell '\) and replaces the output by \(w\, 1^{\rho (r)}\); and \(\mathsf{rd}(\ell _\bot , \ell _0, \ell _1)\) moves to \(\ell _\bot \) if \(p \ge |x|\), and otherwise moves to \(\ell _0\) or \(\ell _1\) according as the symbol \(x_p\) (indexing from \(0\)) is \(0\) or \(1\), and increases the read position to \(p + 1\).
Let \(P\) be a counter program on a binary input \(x\), \(K \in \mathbb {N}\), and \(s\) an abstract state of \(P\) with output \(o\) in which every register holds at most \(K\). Then the output after one step of \(P\) on \(x\) from \(s\) has length at most \(|o| + K + 1\).
Let \(P\) be a counter program over \(R\) registers with labels in \(\Lambda \), \(x\) an input bit string, and \(s\) an abstract state whose read position \(p\) satisfies \(p \le |x|\). Then the read position of the state reached from \(s\) by one step of \(P\) on \(x\) is again at most \(|x|\).
Let \(P\) be a counter program over \(R\) registers with labels in \(\Lambda \), \(x\) an input bit string, \(s\) an abstract state with registers \(\rho \), and \(r\) a register. If \(\rho '\) denotes the registers of the state reached from \(s\) by one step of \(P\) on \(x\), then \(\rho '(r) \le \rho (r) + 1\).
Let \(\Lambda \) be a finite set of labels, \(P\) a counter program over \(R\) registers with labels in \(\Lambda \), \(\ell _0 \in \Lambda \), and \(x\) an input bit string. Let \(C\) be a configuration on input \(x\) of the machine of \(P\) started at \(\ell _0\) which is in some control state \(q\) (not halted), with work tapes \((T_i)\) and work heads \((h_i)\). Then one step of that machine maps \(C\) to the result of applying to \(C\) the action that the compiled transition table of \(P\) assigns to \(q\), the symbol under the input head of \(C\), and the work-tape symbols \((T_i(h_i))_i\).
For \(v \in \mathbb {N}\) let \(1^v\) be the unary tape whose cells \(0, \dots , v-1\) hold \(1\) and whose other cells are blank. Let \(\rho : \{ 0, \dots , R-1\} \to \mathbb {N}\), let \(t\) be an index and \(m \in \mathbb {N}\), and let \(\rho '\) agree with \(\rho \) except that \(\rho '(t) = m\). Then replacing the \(t\)-th entry of the tape vector \((1^{\rho (r)})_r\) by \(1^m\) gives the tape vector \((1^{\rho '(r)})_r\).
Given a list \(\tau = (m_0, \dots , m_{n-1})\) of template micro-operations, labels \(\lambda : \mathbb {N} \to \Lambda \) and a continuation label \(\nu \), the counter-program instruction at index \(i\) is: \(\mathsf{out}(b, \lambda (i+1))\) (print \(b\), go to \(\lambda (i+1)\)) if \(i \lt n\) and \(m_i = \mathsf{out}(b)\); \(\mathsf{pr}(r, \lambda (i+1))\) (print register \(r\) in unary, go to \(\lambda (i+1)\)) if \(i \lt n\) and \(m_i = \mathsf{pr}(r)\); and \(\mathsf{goto}\ \nu \) if \(i \ge n\).
Let \(\Lambda \) be a finite set of labels, \(P\) a counter program over \(R\) registers with labels in \(\Lambda \), and \(\ell _0 \in \Lambda \) a start label. The machine of \(P\) started at \(\ell _0\) is the finite-state multi-tape Turing machine over the bit alphabet with \(R\) work tapes (one per register), whose states are the compiled control states for \(R\) and \(\Lambda \), whose initial state is \(\mathsf{main}(\ell _0)\), and whose transition function is the compiled transition table of \(P\).
Let \(P\) be a counter program over \(R\) registers with labels in \(\Lambda \). Its compiled transition table maps a control state, the symbol under the input head and the symbols under the \(R\) work heads to an action; unless stated otherwise the action writes nothing, moves no head, emits nothing, and touches only the work tape of the register \(r\) involved. In \(\mathsf{main}(\ell )\) it acts according to \(P(\ell )\): \(\mathsf{halt}\) halts; \(\mathsf{goto}(\ell ')\) passes to \(\mathsf{main}(\ell ')\); \(\mathsf{out}(b, \ell ')\) emits \(b\) and passes to \(\mathsf{main}(\ell ')\); \(\mathsf{inc}(r, \ell ')\) writes \(1\) on tape \(r\), moves its head right and passes to \(\mathsf{main}(\ell ')\); \(\mathsf{dec}(r, \ell ')\), \(\mathsf{jz}(r, \ell _0, \ell _+)\) and \(\mathsf{pr}(r, \ell ')\) move the head of tape \(r\) left and pass to \(\mathsf{dec}_2(r, \ell ')\), \(\mathsf{jz}_2(r, \ell _0, \ell _+)\) and \(\mathsf{print}_{\leftarrow }(r, \ell ')\) respectively; and \(\mathsf{rd}(\ell _\bot , \ell _0, \ell _1)\) passes to \(\mathsf{main}(\ell _\bot )\) if the input head reads a blank, and otherwise moves the input head right and passes to \(\mathsf{main}(\ell _0)\) or \(\mathsf{main}(\ell _1)\) according as it reads \(0\) or \(1\). In the auxiliary states, with \(c\) the cell scanned on tape \(r\): \(\mathsf{dec}_2(r, \ell )\) erases \(c\) if it is nonblank and otherwise moves right, then passes to \(\mathsf{main}(\ell )\); \(\mathsf{jz}_2(r, \ell _0, \ell _+)\) moves right and passes to \(\mathsf{main}(\ell _+)\) if \(c\) is nonblank and to \(\mathsf{main}(\ell _0)\) otherwise; \(\mathsf{print}_{\leftarrow }(r, \ell )\) emits \(1\), moves left and stays put in its state if \(c\) is nonblank, and otherwise moves right and passes to \(\mathsf{print}_{\rightarrow }(r, \ell )\); and \(\mathsf{print}_{\rightarrow }(r, \ell )\) moves right and stays in its state if \(c\) is nonblank, and otherwise passes to \(\mathsf{main}(\ell )\) without moving.
For \(R\) registers and a set \(\Lambda \) of labels, the control states of the Turing machine compiled from a counter program are of five kinds, where \(r\) is a register and \(\ell , \ell _0, \ell _+\) are labels: \(\mathsf{main}(\ell )\), about to execute the instruction at label \(\ell \); \(\mathsf{dec}_2(r, \ell )\), the second phase of a decrement of register \(r\); \(\mathsf{jz}_2(r, \ell _0, \ell _+)\), the second phase of a zero test of register \(r\); \(\mathsf{print}_{\leftarrow }(r, \ell )\), printing register \(r\) while walking left over its cells; and \(\mathsf{print}_{\rightarrow }(r, \ell )\), walking back right after such a print. When \(\Lambda \) is finite, so is the set of control states.
The work-tape contents determined by a finite binary word \(b\): for an integer position \(z\), cell \(z\) holds the \(z\)-th bit of \(b\) when \(0 \le z \lt |b|\), and is blank at every other position (in particular at all negative positions).
Let \(p\) and \(b\) be binary words, lay their concatenation on the nonnegative cells of a tape, and let \(\beta \) be a bit. Overwriting cell \(|p|\) with \(\beta \) yields exactly the tape of the word obtained by following \(p\) with \(\beta \) and then the tail of \(b\) (all bits of \(b\) after its first). In particular, when \(b\) is empty the stored word is extended by one cell; all other cells, including the negative ones, are unchanged.
A finite Turing machine over the binary alphabet with one work tape and four states: count, carry, rewind, and emit. In the count state it performs a carry transition while also advancing the input head, or switches to the emit state when the input is exhausted. In the carry state it propagates an increment along the work tape (clearing \(1\)s rightward until it writes a \(1\) over the first \(0\) or blank). In the rewind state it moves the work head left until it finds the blank at cell \(-1\), then steps right to cell \(0\) and re-enters the count state. In the emit state it copies the work-tape word to the output, one bit per step moving right, and halts at the first blank.
Given an optional symbol of the extended simulation alphabet, its decoding is a pair of payloads: a data-cell symbol carrying a current payload and a cached left-neighbor payload decodes to exactly that pair, while every other symbol, including the untouched physical blank, decodes to the pair of blank payloads.
Let \(W\) and \(M\) be binary machines and \(a\) a padding constant. Given a base schedule configuration, a schedule phase \(q\), unary and guide head positions \(u, g \in \mathbb {Z}\), a finite data register \(d\), contents for the \(k+1\) data tapes, and an output list, the associated configuration of the logical simulator is the decorated form of the macrostep configuration of the base at phase \(q\) with heads \(u\) and \(g\), carrying \(d\), the data tapes, and the output.
Suppose the base’s guide tape is the fixed guide of radius \(R\), let \(f\) assign a payload vector to every coordinate, and suppose the current payloads of the data tapes decode to \(f\) while the cached left neighbors equal \(f(z-1)\) throughout \([-R, R]\). For every \(n \le 2R + 1\), running the simulator \(n\) steps from the backward phase at \(R\) with neighbor registers \(f(R+1)\) stays in the backward phase, reaching position \(R - n\) with neighbor registers \(f(R-n+1)\), on tapes whose current payload on lane \(i\) at coordinate \(z\) is the shifted value \(f(z + m_i)\) for visited coordinates \(R - n \lt z \le R\) — where \(m_i\) is the virtual head direction of lane \(i\) selected from the saved state and reads — and the unshifted value \(f(z)\) elsewhere, while the cached left neighbors remain \(f(z-1)\) for \(-R \le z \le R - n\).
Suppose the base’s guide tape is the fixed guide of radius \(R\) and the simulator is in the backward phase with guide head at \(g \neq -R - 1\). One step moves the guide head to \(g - 1\), replaces each neighbor register by the current payload decoded at \(g\), and rewrites cell \(g\) of each data lane to a data cell with blank cached neighbor whose current payload is chosen by that lane’s virtual head direction: the cached left neighbor for a leftward move, the unchanged current payload for a stationary move, and the carried register for a rightward move.
Suppose the base’s guide tape is the fixed guide of radius \(R\). Running the simulator exactly \(R + 1\) steps from the return phase with guide head at \(-R\) and unary head at \(u\) reaches the counter-test phase with unary head \(u + 1\) and guide head at the origin; the simulated state is replaced by the successor state of the source action selected from the saved state and reads, while the answer bit, reads, neighbor registers, data tapes, and output are unchanged.
Suppose the base’s unary tape holds a unit marker at \(u\). One simulator step from the counter-test phase saves the current payloads decoded at the origin as the new reads, updates the answer register to the bit emitted by the source action selected from the saved state and those reads (keeping the old answer when nothing is emitted), performs that action’s optional work-lane writes at the origin, clears the neighbor registers to blank, and enters the leftward-seek phase; no output is emitted.
If the base’s unary tape does not hold a unit marker at \(u\), then for every data register and all data tapes, one simulator step from the counter-test phase with empty output halts, and its output is the single bit symbol carrying the saved answer bit of the register.
Suppose the base’s guide tape is the fixed guide of radius \(R\), let \(f\) assign a payload vector to every integer coordinate, and suppose every current payload of the data tapes decodes to \(f\). For every \(n \le 2R + 1\), running the simulator \(n\) steps from the forward phase at \(-R\) with neighbor registers \(f(-R-1)\) stays in the forward phase, reaching position \(-R+n\) with neighbor registers \(f(-R+n-1)\), on tapes whose current payloads still decode to \(f\) and whose cached left-neighbor payloads equal \(f(z-1)\) at every visited coordinate \(-R \le z \lt -R + n\).
Suppose the base’s guide tape is the fixed guide of radius \(R\) and the simulator is in the forward phase with guide head at \(g \neq R + 1\). One step keeps the saved state, answer bit, and reads, moves the guide head to \(g+1\), replaces each neighbor register by the current payload decoded at \(g\), and rewrites cell \(g\) of each data tape to a data cell carrying its unchanged current payload with the old neighbor register cached; no output is emitted.
Suppose the base’s guide tape is the fixed guide of radius \(R\). From the forward phase with guide head at \(R + 1\), one simulator step enters the backward phase at position \(R\), resets every neighbor register to the blank payload, and leaves the saved state, answer bit, reads, data tapes, and output unchanged.
Let the base configuration’s unary tape have \(B\) markers with \(|x| + 1 \le B\) and its guide tape have radius \(3B\), and let the data tapes decode the head-centered payload vectors of the initial configuration of the source \(M\) on \(x\). Then for every \(j \le B\) there are reads, neighbor registers, and data tapes such that \(j\, (18B + 8)\) simulator steps from the first counter-test configuration reach the counter-test phase with unary head \(j + 1\), carrying the state and last-emitted bit of the \(j\)-step run of \(M\) on \(x\), with data tapes decoding that run’s head-centered payload vectors and no output emitted.
Suppose the base’s guide tape is the fixed guide of radius \(R\). Running the simulator exactly \(R + 2\) steps from the leftward-seek phase with guide head at the origin and blank neighbor registers reaches the forward phase at \(-R\) with blank neighbor registers, leaving the saved state, answer bit, reads, data tapes, and output unchanged.
Suppose the base’s guide tape is the fixed guide of radius \(R\) and its unary tape holds a unit marker at \(u\), and let \(c\) be a configuration of the source \(M\) whose head-centered payload vectors are decoded by the data tapes and are blank outside \([-R, R]\), as are those of its one-step successor. Then exactly \(6R + 8\) simulator steps from the counter-test phase carrying the state and last-emitted bit of \(c\) reach the counter-test phase at unary head \(u + 1\) carrying the state and last-emitted bit of the successor of \(c\), with data tapes decoding the successor’s head-centered payload vectors and an unchanged output stream.
Let \(W\) and \(M\) be binary machines and \(a\) a padding constant. One step of the logical simulator on a macrostep configuration with data equals the decorated form of one schedule-controller step, where the new finite data register, the single cell written under the guide head on each data tape, and the appended emission are exactly those returned by the data transducer evaluated at the current phase, register, input read, schedule tape reads, and the data cells under the guide head.
Suppose the base’s guide tape is the fixed guide of radius \(R\), the current payloads of the data tapes decode to \(f\), and the boundary values \(f(-R-1)\) and \(f(R+1)\) are blank. Then running the simulator exactly \(6R + 7\) steps from the leftward-seek phase at the origin with unary head \(u\) and blank neighbor registers reaches the counter-test phase with unary head \(u + 1\), simulated state equal to the successor state of the source action selected from the saved state and reads, and neighbor registers \(f(-R)\), on tapes whose current payload on lane \(i\) at coordinate \(z\) is \(f(z + m_i)\) for \(-R \le z \le R\) and \(f(z)\) outside, where \(m_i\) is that lane’s selected virtual head direction.
Let \(M\) be a binary machine and \(c\) a configuration of \(M\), and suppose the current payloads of the data tapes decode to the head-centered payload vectors of \(c\). After writing at the origin of each work lane the symbol optionally written by the totalized action of \(M\) at \(c\) (with blank cached neighbor, the input lane untouched), the current payloads decode to the written payload vectors of \(c\) under that action.
Given binary machines \(W\) (the clock witness) and \(M\) (the source, with \(k\) work tapes) and a padding constant \(a\), the logical simulator is the machine over the extended simulation alphabet obtained by decorating the length-only schedule controller built from \(W\) and \(a\) with \(k+1\) additional data tapes tracked at the guide-tape head, starting from the initial data state of \(M\) and driven by the data transducer of \(W\), \(M\), and \(a\).
Let \(W\) compute on every input \(x\) the binary expansion of \(T(|x|)\) within \(b\, (T(|x|)+1)\) steps, suppose \(n \le T(n)\) for all \(n\), and let \(M\) decide the language \(L\) within time \(a\, T(n)\). Then for every word \(x\) there is a time \(t\) within which the logical simulator, on the embedded input, computes the one-symbol output consisting of the bit symbol of the indicator of \(L\) at \(x\).
Let \(W\) compute on every input \(x\) the binary expansion of \(T(|x|)\) within \(b\, (T(|x|)+1)\) steps, and suppose \(n \le T(n)\) for all \(n\). Then for every word \(x\) there are a time \(t\), a base schedule configuration whose unary tape has \((a+1)(T(|x|)+1)\) markers and whose guide tape has radius \(3(a+1)(T(|x|)+1)\), and data tapes decoding the head-centered payload vectors of the initial configuration of the source \(M\) on \(x\), such that the \(t\)-step run of the logical simulator on the embedded input is the macrostep configuration with data at the counter-test phase, unary head \(1\), guide head at the origin, initial data register, and empty output.
Let \(c\) be a configuration of a \(k\)-tape machine with state set \(S\) on an input word \(x\) over an alphabet \(A\), let \(h \in \{ 0, \dots , k-1\} \) be a designated head index, let \(d\) be an element of a set \(D\) of data states, let \(l\) tape contents \(\mathbb {Z}\to A \cup \{ \text{blank}\} \) be given, and let an output word over \(A\) be given. The associated decorated configuration is the configuration of a \((k+l)\)-tape machine with state set \(S \times D\) on the same input \(x\) whose control state is the state of \(c\) paired with \(d\) (halted stays halted), whose input head agrees with \(c\), whose first \(k\) work tapes and head positions are those of \(c\), whose last \(l\) (data) tapes carry the given contents with every data head at the position of the designated head \(h\) of \(c\), and whose output is the given word, maintained separately from the silent schedule.
Let \(P\) be a machine over an alphabet \(A\), decorated by a data transducer with \(l\) extra tapes, designated head index \(h\), a finite data-state set with initial element \(d_0\), and a visit function. For every input word \(x\), the initial configuration of the decorated machine on \(x\) equals the decorated configuration built from the initial configuration of \(P\) on \(x\), the initial data state \(d_0\), entirely blank extra tapes, and the empty output: the data tapes begin blank and aligned with the schedule’s origin.
Let \(P\) be a machine over an alphabet \(A\), decorated by a data transducer with \(l\) extra tapes, designated head index \(h\), initial data state, and visit function. Let \(c\) be a configuration of \(P\) on an input word \(x\) whose control state is a live state \(q\), let \(d\) be a data state, and fix contents for the \(l\) data tapes and an output word. Write \(v\) for the value of the visit function on \(q\), \(d\), the input symbol scanned by \(c\), the \(k\) work-tape reads of \(c\), and the data-tape symbols at the position of the designated head \(h\); thus \(v\) consists of a new data state, one optional write per data tape, and an optional emitted symbol. Then one step of the decorated machine from the decorated configuration equals the decorated configuration built from one step of \(P\) from \(c\), the new data state, each data tape updated by its optional write at the position of the designated head, and the output word extended by the emitted symbol when one is present. In particular the decorated transition consists of exactly the prescribed schedule step together with the local data visit.
Let \(P\) be a machine over an alphabet \(A\) with \(k\) work tapes, let \(l \in \mathbb {N}\), let a designated head index \(h \in \{ 0, \dots , k-1\} \) be given, and let \(D\) be a finite set of data states with a distinguished initial element. A visit function maps a state of \(P\), a data state, the scanned input symbol, the \(k\) schedule reads, and the \(l\) data reads to a new data state, a write for each of the \(l\) data tapes, and an optional emitted symbol. The decorated machine has \(k + l\) work tapes and state set \(P\)’s states paired with \(D\). On the first \(k\) tapes it performs exactly \(P\)’s transitions; each of the \(l\) extra tapes performs the write chosen by the visit function and moves exactly as the designated head \(h\) does; the emitted output is the one chosen by the visit function; and the decorated machine halts exactly when \(P\) halts.
Let a machine \(P\) over an alphabet \(A\) be decorated by any data transducer with \(l\) extra tapes and designated head \(h\). For every input word \(x\), every time \(t \in \mathbb {N}\), and every extra-tape index \(i \lt l\), the head position of the \(i\)-th data tape of the decorated machine after \(t\) steps equals the head position of the designated schedule head \(h\) at the same time — including at all times after halting. Consequently no data-selected write can alter a head trajectory.
Let \(P\) be a machine over an alphabet \(A\) whose transition table never emits an output symbol, and suppose \(P\) is oblivious. Then for every choice of data transducer — number of extra tapes, designated head, finite data-state set with initial element, and visit function — the decorated machine is oblivious. The data states, emitted answers, and data-selected writes may depend arbitrarily on the actual input; none of them can choose a head movement or terminate the schedule early.
Let \(P\) be a machine over an alphabet \(A\) whose transition table never emits an output symbol, decorated by any data transducer. For every input word \(x\) and every time \(t \in \mathbb {N}\), the schedule projection of the decorated machine’s configuration after \(t\) steps from its initial configuration on \(x\) equals the configuration of \(P\) after \(t\) steps from its initial configuration on \(x\).
Let \(P\) be a machine over an alphabet \(A\) whose transition table never emits an output symbol, and consider its decoration by any data transducer with \(l\) extra tapes, designated head, initial data state, and visit function. For every configuration \(c\) of the decorated machine, the schedule projection of one step of the decorated machine at \(c\) equals one step of \(P\) at the schedule projection of \(c\). This is an equality of entire configurations, not merely a correspondence at macrostep boundaries.
For a time bound \(T : \mathbb {N} \to \mathbb {N}\), the class \(\mathrm{DTIME}(T)\) consists of all languages \(L\) over the binary alphabet for which there exist a constant \(c \in \mathbb {N}\) and a finite multi-tape Turing machine \(M\) over the binary alphabet such that on every input \(x\) the machine \(M\) halts within \(c \cdot T(|x|)\) steps and outputs \(1\) if \(x \in L\) and \(0\) otherwise.
Let \(T_1, T_2 : \mathbb {N} \to \mathbb {N}\) be time bounds with \(T_1(n) \le T_2(n)\) for every \(n\). Then \(\mathrm{DTIME}(T_1) \subseteq \mathrm{DTIME}(T_2)\): every language over the binary alphabet decided by some finite multi-tape machine within \(c \cdot T_1(n)\) steps on inputs of length \(n\), for some constant \(c\), is also decided by such a machine within \(c \cdot T_2(n)\) steps.
Let \(T : \mathbb {N} \to \mathbb {N}\) be a time bound with \(T(n) = 0\) for at least one input length \(n\). Then \(\mathrm{DTIME}(T) = \emptyset \): no language over the binary alphabet is decidable by a finite multi-tape machine within \(c \cdot T\) steps for any constant \(c\), because on an input of length \(n\) the allotted budget is \(c \cdot T(n) = 0\) and no machine halts in zero steps.
For every natural number \(c\), the class \(\mathrm{DTIME}(n \mapsto n^c + 1)\) — the languages over \(\{ 0,1\} \) decided by some deterministic finite Turing machine within \(a \cdot (n^c + 1)\) steps on inputs of length \(n\), for some constant \(a\) — is contained in \(\mathsf{P}\).
For every time bound \(T : \mathbb {N} \to \mathbb {N}\), \(\mathrm{DTIME}(T) \subseteq \mathrm{NTIME}(T)\). That is, if a language \(L\) of binary strings is decided by a finite deterministic multi-tape Turing machine that, for some constant \(c\), halts on every input \(x\) within \(c \cdot T(|x|)\) steps with output \(1\) if \(x \in L\) and \(0\) otherwise, then \(L\) is decided by a finite nondeterministic multi-tape machine within \(c \cdot T(|x|)\) steps on every input \(x\).
Let \(N\) be a nondeterministic Turing machine over the binary alphabet, \(C, c\) natural numbers, and \(x, u\) binary strings with \(|u| = C \cdot 2^{(|x|+1)^c}\). Then \(x\, u\) lies in the exponential choice-word verifier language of \(N\) with parameters \(C, c\) if and only if the run of \(N\) on \(x\) under \(u\) has halted with output the one-bit string \(1\).
Let \(N\) be a nondeterministic Turing machine deciding a language \(L\) within time \(a \cdot 2^{n^c}\) on inputs of length \(n\). For every binary string \(x\): \(x \in L\) if and only if there is a string \(u\) with \(|u| = a \cdot 2^{(|x|+1)^c}\) such that \(x\, u\) lies in the exponential choice-word verifier language of \(N\) with parameters \(a, c\).
Let \(N\) be a nondeterministic Turing machine and \(C, c\) natural numbers. If the exponential split search fails at the length of a string \(y\), then \(y\) is not in the exponential choice-word verifier language of \(N\) with parameters \(C, c\).
For all natural numbers \(C\) and \(c\) there exist a deterministic machine and a constant \(A\) computing, on every input \(x\), the little-endian binary expansion of \(C \cdot 2^{(|x|+1)^c}\) within \(A\, (|x|+1)^{c+1}\) steps.
For all natural numbers \(C, c\) and every word \(w\), searching the indices \(i \le |w|\) for the first round state (the \(i\)-fold iterate from the empty state) that passes the round test returns exactly the result of the exponential split search at length \(|w|\).
For natural numbers \(C, c\) and a word \(w\): taking the first accepted index of the bounded orbit search and emitting the encoded pair of the corresponding prefix and suffix of \(w\) (or the empty word when the search is exhausted) yields exactly the word of the exponential split emitter on \(w\).
For every fixed word \(w\), every input \(x\) and every \(t \le |x|\): after \(t\) steps from the initial configuration, the zero-prefixing emitter for \(w\) is still in its scanner state, its input head is at position \(t + 1\), and its output so far is \(t\) zero bits.
Let \(N\) be a nondeterministic Turing machine and \(C, c\) natural numbers. For every binary string \(y\), the pair-loader machine of \(N\), run on the word emitted by the exponential split emitter for \(y\), halts within \(3(2|y|+3)\) steps, and its output is the indicator bit of membership of \(y\) in the exponential choice-word verifier language of \(N\).
Fix natural numbers \(C, c\) and suppose a deterministic loop-body machine with a designated anchor state satisfies two timed contracts within \(A(|w|+1)^{r+1}\) steps on every input \(w\): a startup contract that first reaches the anchor with clean scratch, and a round contract that, from any anchored search state, either halts with the encoded pair payload when the round test accepts, or returns to the anchor with the stepped search state. Then some deterministic machine computes the exponential split-emitter function within \(B(n+1)^{r+2}\) steps for a constant \(B\).
Let \(N\) be a nondeterministic Turing machine and \(C, c\) natural numbers. If some deterministic machine computes the exponential split-emitter function within time \(A(n+1)^{r+1}\) for constants \(A, r\), then the exponential choice-word verifier language of \(N\) with parameters \(C, c\) belongs to \(\mathrm{P}\).
For all binary strings \(u, v\) and every index \(j\), the length-\((j+1)\) prefixes of \(u\) and \(v\) agree if and only if their length-\(j\) prefixes agree and their optional bits at index \(j\) agree, a missing bit counting as different from any present bit.
Given a nondeterministic Turing machine \(N\) over the binary alphabet and natural numbers \(C, c\), the exponential choice-word verifier language consists of all strings \(x\, u\) (concatenation) with \(|u| = C\cdot 2^{(|x|+1)^c}\) such that the run of \(N\) on \(x\) under choice word \(u\) has halted with output the one-bit string \(1\).
For natural numbers \(C, c\) and a binary string \(y\), the emitted word is the self-delimiting encoding of the pair formed by the first \(i\) bits of \(y\) and the remaining bits of \(y\), when the exponential split search returns \(i\) at length \(|y|\); it is the empty word on failure.
Let \(M\) be a binary multi-tape Turing machine with finite state set \(Q\), \(k\) work tapes, an append-only output tape and transition function \(\delta \), and let \((q, a, b)\) be a snapshot of \(M\) (a state \(q \in Q \cup \{ \bot \} \) with \(\bot \) meaning halted, an input symbol \(a\), and work-tape symbols \(b\)). The emitted symbol of the snapshot is either a bit or nothing: it is nothing if \(q = \bot \), and otherwise it is the optional output bit that the transition \(\delta (q, a, b)\) appends to the output tape.
Given an acceptance function on natural-number ranks, a start rank \(i\), and a count: the Boolean that is true exactly when one of the \(\mathrm{count}\) consecutive ranks starting at \(i\) accepts, evaluated by testing the ranks in order.
Let \(V\) be a language over the binary alphabet, \(x\) a string, and \(w\) a width. Testing all \(2^w\) ranks with the acceptance function “the indicator of \(V\) accepts \(x\) concatenated with the width-\(w\) binary word of the rank” returns true if and only if there exists \(u\) with \(|u| = w\) and \(x \mathbin {+\! \! +} u \in V\).
For every width \(w\) and every predicate \(p\) on boolean words: there exists a word \(u\) with \(|u| = w\) satisfying \(p\) if and only if there exists a rank \(i \lt 2^w\) such that \(p\) holds of the width-\(w\) binary word of \(i\).
For every candidate word \(s\), after one step of the enumerator body (active or stopped) from the two-ended configuration in the verdict state with tape zero \(s\) and the captured word consisting of a single true bit, the machine is halted and its output is exactly the single bit true.
Let \(B\) be a source machine with at least one work tape that, started in state \(q\) at the seam carrying candidate \(s\) on input \(x\), halts within \(T\) steps with output \(y\). Then the enumerator body, from the two-ended configuration embedding the clean call of \(q\) in a given mode, reaches within \(3T + |x| + |y| + 8\) steps the two-ended configuration in that mode’s administrative continuation state, with candidate \(s\) preserved, \(y\) on the capture tape, all scratch blank, both endpoint heads at zero, and empty physical output.
One step of the enumerator body from the increment-check state, at the two-ended configuration with candidate \(s\) and captured increment result \(v\), preserves all tapes and transfers control to the anchor when \(v\) is empty (overflow) and to the result-copy state otherwise.
For every input \(x\) and candidate word \(s\), one step of the active enumerator body from the anchor state over the canonical seam carrying \(s\) on tape zero yields the two-ended configuration that enters the verification call (the source’s verifier entry state in verification mode), with tape zero \(s\), a blank capture tape, and both endpoint heads at zero; the step writes nothing, moves no head, and emits nothing.
Let \(B\) be a source machine with at least one work tape that, started in the increment entry state at the seam with candidate \(s\), halts within \(T\) steps emitting the fixed-width increment’s output word. Then the enumerator body, from the two-ended configuration embedding the increment call, reaches within \(3T + |x| + 3|s| + 11\) steps the anchor over the seam carrying the stalled increment of \(s\): the successor word of the same width, or \(s\) itself on overflow.
For every candidate word \(s\), one step of the enumerator body (active or stopped) from the two-ended configuration in the verdict state with tape zero \(s\) and a captured single false bit yields the two-ended configuration that enters the increment call (the source’s increment entry state in increment mode), with tape zero \(s\), the captured verdict erased so that the capture tape is blank, and both endpoint heads at zero.
Let \(B\) be a source machine with at least one work tape whose verification call on candidate \(s\) halts within \(T_v\) steps with single-bit verdict \(b\), and whose increment call halts within \(T_i\) steps emitting the fixed-width increment output. Then there is a time \(t\) with \(0 \lt t \le 3T_v + 3T_i + 2|x| + 3|s| + 22\) such that the active enumerator body, started at the anchor over the seam with candidate \(s\), does not visit the anchor at any strictly intermediate positive time, and at time \(t\) either halts with output true (when \(b\) is true) or stands at the anchor over the seam carrying the stalled increment of \(s\) (when \(b\) is false).
Let \(B\) be a source machine with at least one work tape whose verification call on candidate \(s\) (from the verifier entry state at the seam) halts within \(T_v\) steps with single-bit output \(b\), and whose increment call halts within \(T_i\) steps emitting the fixed-width increment output. Then from the configuration just after departure, within \(3T_v + 3T_i + 2|x| + 3|s| + 21\) steps the enumerator body either halts with output true (when \(b\) is true) or reaches the anchor over the seam carrying the stalled increment of \(s\) (when \(b\) is false).
Let \(B\) be a source machine with at least one work tape that, started at the empty seam on input \(x\), halts within \(T\) steps emitting \(w\) consecutive true marks. Then within \(3T + |x| + 3w + 10\) steps from its initial configuration on \(x\), the enumerator body reaches the anchor state over the canonical seam carrying the all-false candidate of width \(w\).
Let \(B\) be a source machine with at least one work tape that, from the empty seam on input \(x\), halts within \(T\) steps emitting \(w\) true marks. Then within \(3T + |x| + 3w + 10\) steps the active enumerator body reaches the anchor over the all-false width-\(w\) candidate seam, without visiting the anchor at any earlier time.
Let \(B\) be a finite source machine with designated verification and increment entry states. The enumerator body built from \(B\) is the machine that issues clean calls of \(B\) (with halts redirected to live administrative states) for initialization, verification, and increment, coordinated by seven administrative states: an anchor, a seed-copy phase, an unused reserve, a verdict test, an increment check, a result copy, and a synchronized rewind. Its stopped variant agrees with the active body except that the anchor is a silent self-loop.
For any displacements \(m, d\) and control states \(q, q'\), applying the two-head administrative action with displacements \(m, d\) and next state \(q'\) to an administrative configuration of the clean wrapper in state \(q\) yields the administrative configuration with the same tapes, retained word and empty output, in state \(q'\), with the native input head moved by \(m\) and the capture head moved by \(d\).
Given displacements \(m\) and \(d\) and a next control state, the two-head administrative action is the action on \(L + 1\) work tapes that writes nothing, emits nothing, moves the native input head by \(m\) and the head of the last (capture) tape by \(d\), leaves the heads of the first \(L\) work tapes in place, and selects the given next control state.
For every \(j \le |y|\): from the administrative configuration of the clean wrapper with retained word \(y\) and capture head just left of cell \(j\), exactly \(j + 1\) steps return the capture head to zero without erasing \(y\), advancing control to the final administrative state.
If a source machine \(M\) is live for \(t\) steps from a configuration \(c_0\), then the clean wrapper of \(M\), run for \(t\) steps from the captured logged start of \(c_0\), reaches the captured logged configuration of \(M\)’s time-\(t\) configuration with the \(t\)-step history; any halting emission of \(M\) is retained on the capture tape.
If a machine \(M\) halts within \(T\) steps from configuration \(c_0\) on input \(x\) with output \(y\), then within \(3T + |x| + |y| + 8\) steps the clean wrapper of \(M\), started at the captured logged start of \(c_0\), halts in the administrative configuration with \(c_0\)’s work fields restored, all histories blank, \(y\) retained on the last tape, the input head at cell one, the capture head at zero, and no physical output.
For a machine \(B\) with at least one work tape: the clean wrapper’s captured entry configuration at the prepared seam carrying \(s\) on tape zero equals the two-ended administrative configuration with tape zero \(s\), blank capture, and both endpoint heads at zero.
For a machine \(B\) with at least one work tape: the clean wrapper’s completed administrative configuration over the prepared seam carrying \(s\), with retained output \(y\), equals the two-ended configuration with tape zero \(s\), capture tape \(y\), and both endpoint heads at zero.
If the source \(M\) is live for \(t\) steps from \(c_0\), then within \(2t + 1\) steps the clean wrapper’s captured restoration, started at the checkpoint of the time-\(t\) configuration with its history, reaches the administrative configuration in which the original source work fields of \(c_0\) are restored, the time-\(t\) output is retained on the last tape, and the capture head stands at the end of that word.
If the source configuration \(c\) is halted, then one silent step of the clean wrapper from the captured logged configuration of \(c\) with history \(h\) parks every history head on the last entry and enters the captured restoration phase at the checkpoint for \(c\) and \(h\), retaining \(c\)’s output on the capture tape.
An administrative configuration of the clean-call wrapper: the source block holds the restored work fields of a given source configuration with blank histories, the last tape holds a retained word \(y\), and the control state, native input-head position, and capture-head position are parameters. The output is empty.
The clean subroutine built from a finite machine \(M\): it simulates the logging version of \(M\) with the output captured on one extra final tape, then runs the captured restoration of all source work, rewinds the native input head and the captured word, and halts silently. The retained output word lives on the last work tape; no physical output is emitted.
Suppose the candidate word \(s\) decomposes as \(\mathit{pre} \mathbin {+\! \! +} \mathit{rest}\), and let \(\mathit{out}\) be a word and \(p\) an input-head position. After exactly \(|\mathit{rest}| + 1\) steps from the candidate-phase configuration of the assembly machine with input head at \(p\), candidate head at position \(|\mathit{pre}|\), and output \(\mathit{out} \mathbin {+\! \! +} \mathit{pre}\), the machine is halted with the candidate tape unchanged, its head at position \(|s|\), the input head still at \(p\), and output \(\mathit{out} \mathbin {+\! \! +} s\); an empty remainder takes the single final step.
Suppose the input \(x\) decomposes as \(\mathit{pre} \mathbin {+\! \! +} \mathit{rest}\), and let \(s\) be a candidate word. After exactly \(|\mathit{rest}| + 1\) steps from the input-phase configuration of the assembly machine with the input head on the first symbol of \(\mathit{rest}\), candidate \(s\) with its head at zero, and output \(\mathit{pre}\), the machine is in the candidate phase with the input head at the right boundary, the candidate tape and its head at zero unchanged, and output \(x\).
For all binary words \(x\) and \(s\), the assembly machine started in its input phase on input \(x\), with \(s\) on its work tape and all heads at their starting positions, is halted after exactly \(|x| + |s| + 2\) steps with output \(x \mathbin {+\! \! +} s\), its input head at the right boundary, the candidate tape unchanged, and the candidate head at position \(|s|\).
An assembly configuration is the configuration of the assembly machine on input \(x\) determined by a control state, an input-head position, a candidate word \(s\) held fixed on the single work tape with its head at a given cell, and the output prefix emitted so far.
For a machine with copy and rewind states as in the two preceding scans: if the captured word \(v\) is at least as long as the old tape-zero word, then from a two-ended configuration with both heads at zero, exactly \(2|v| + 2\) steps replace tape zero entirely by \(f\) mapped over \(v\), blank the capture tape, and return both heads to zero in the arrival state.
Consider a machine with at least two work tapes whose copy state, at a two-ended configuration, overwrites the tape-zero cell with \(f\) applied to the captured symbol, blanks the captured cell, and advances both heads, exiting left into a rewind state at the captured right blank. Starting with tape zero holding \(\mathit{pre} \mathbin {+\! \! +} \mathit{old}\), the capture tape holding \(\mathit{rest}\) beyond \(|\mathit{pre}|\) already-erased cells, and both heads at \(|\mathit{pre}|\): after \(|\mathit{rest}| + 1\) steps the machine is in the rewind state, tape zero holds \(\mathit{pre}\), then \(f\) mapped over \(\mathit{rest}\), then the remainder of \(\mathit{old}\) beyond position \(|\mathit{rest}|\), the capture tape is blank, and both heads stand at \(|\mathit{pre}| + |\mathit{rest}| - 1\).
Let \(L \gt 0\), and consider a two-ended configuration on \(L + 1\) work tapes whose tape zero holds \(u\) and whose last tape holds \(v\), with common head position \(h\). Applying the two-ended action with optional writes \(a\), \(b\), displacement \(d\) and next state \(q'\) yields the two-ended configuration in state \(q'\) whose tape zero is \(u\) with cell \(h\) overwritten by \(a\) (if \(a\) writes), whose last tape is \(v\) with cell \(h\) overwritten by \(b\) (if \(b\) writes), and whose common head position is \(h + d\); the input head and the empty output are unchanged.
Given optional writes \(a\) and \(b\), a displacement \(d\), and a next control state, the two-ended action is the action on \(L + 1\) work tapes that performs write \(a\) on tape zero and write \(b\) on the last tape, moves both of these heads by \(d\), leaves every intervening tape unwritten and stationary, keeps the input head fixed, and emits nothing.
Let two transition tables agree away from an anchor state, which the second (stopped) table turns into a silent self-loop. If the stopped run from \(c\) stands at the anchor at time \(T\), then there is a time \(t \le T\) such that the active run never visits the anchor before \(t\) and its time-\(t\) configuration equals the stopped run’s time-\(T\) configuration.
Fix constants \(C, c\), a budget function \(G\), a language \(V\) over the binary alphabet, and a finite machine (the body) with a designated anchor state satisfying: (i) from its initial configuration on any input \(x\) it reaches, within \(G(|x|)\) steps and with no earlier anchor visit, the anchor over the seam carrying the all-false candidate of width \(w = C(|x|+1)^c\); and (ii) from the anchor over any candidate \(s\) of width \(w\), within some positive time at most \(G(|x|)\) and with no interior anchor visit, it halts with output true when \(x \mathbin {+\! \! +} s \in V\) and otherwise returns to the anchor over the stalled fixed-width increment of \(s\). Then there exist a constant \(b\) and a machine \(E\) such that for every input \(x\) there is a rank-indexed family of configurations with: \(E\) reaches the rank-\(0\) configuration from its initial configuration within \(b(G(|x|) + (|x|+1)^{c+1} + 1)\) steps; the rank-\(2^w\) configuration is halted with single-bit output false; and from each rank \(i \lt 2^w\), within the same budget, \(E\) either halts with output true (when \(x\) concatenated with the width-\(w\) binary word of \(i\) lies in \(V\)) or reaches rank \(i + 1\).
Let two transition tables agree away from an anchor state, which the second (stopped) table turns into a silent self-loop. If the stopped run from \(c\) is halted at time \(T\), then the active run from \(c\) equals the stopped run at time \(T\), and the active run never visits the anchor at any time up to \(T\).
For every boolean word \(s\), the fixed-width little-endian binary increment of \(s\) computed by the Turing-machine catalog’s counter equals the one computed by the certificate-enumeration counter: both return the same-length incremented word, or both fail exactly on overflow (including on the empty word).
Suppose a machine \(I\) computes, within time \(j(n+1)\) on inputs of length \(n\), the output word of the fixed-width binary increment (the successor word, or the empty word on overflow). Then from the candidate seam in the candidate phase with candidate \(s\) on input \(x\), the buffered composition of the assembly machine with \(I\) halts within \((j+3)(|s|+1)\) steps and outputs the increment result of \(s\), which is empty exactly on overflow.
For every boolean word \(s\), exactly one of the following holds for the fixed-width increment’s output word: it is empty (overflow) and the stalled step leaves \(s\) unchanged; or it is nonempty, has length exactly \(|s|\), and the stalled step equals it.
Suppose a host machine simulates a source machine with at least one work tape by padded transitions on an initial tape block. If the source, started in state \(q\) at the seam carrying \(s\), is halted at time \(T\) with output \(y\), then the host, started in the embedded state at its own seam carrying \(s\), is halted at time \(T\) with the same output \(y\).
Suppose a host machine simulates a source machine by padded transitions on an initial tape block. If the source computes \(y\) from input \(x\) within \(T\) steps, then the host, started in the embedded initial state at the empty seam on \(x\), is halted at time \(T\) with output \(y\); this holds even for a source with no work tapes.
For any configuration \(c\) with history \(h\) and any action \(a\) of the source machine: applying the logged version of \(a\) to the logged configuration of \(c\) with history \(h\) yields the logged configuration of \(a\) applied to \(c\), with history \(h\) extended by the single entry recording the overwritten symbols and head moves of \(a\). This includes halting and output-emitting actions.
For every finite machine \(M\), start configuration \(c_0\) (possibly prepared, not necessarily initial), and time \(t\): running the logging machine of \(M\) for \(t\) steps from the logged configuration of \(c_0\) with empty history yields the logged configuration of \(M\)’s time-\(t\) configuration together with the history of the first \(t\) steps. In particular the source’s input, output, and halting time are unchanged by the instrumentation.
For a machine \(B\) with at least one work tape: the logged configuration, with empty history, of the canonical seam carrying a word \(s\) on tape zero equals the canonical seam of the logging machine’s larger tape bank, the added tapes being blank and stationary.
Given the symbols currently under the \(k\) work heads and an action of a \(k\)-tape machine, the corresponding action of a machine with \(k + (1 + (k + k))\) work tapes: it performs the original action on the first \(k\) tapes while advancing a unary clock tape and recording, on \(2k\) history tapes, the overwritten symbols and the encoded head moves. Output and control are otherwise unchanged.
The configuration of the logging machine corresponding to a source configuration \(c\) and a finite history list \(h\): the source fields occupy the first block of tapes, the clock tape holds \(|h|\) true marks, the history tapes hold the recorded old symbols and encoded moves as sparse tapes, and every history head sits one cell past the recorded entries.
For a finite machine \(M\) with \(k\) work tapes, the logging machine of \(M\) is the machine with the same finite state set and initial state and with \(k + (1 + (k + k))\) work tapes, every transition of which performs \(M\)’s transition on the first \(k\) tapes while logging a clock mark, the overwritten symbols, and the head moves on the extra tapes.
For every width \(w\) and every \(i \lt 2^w\), iterating the stalled fixed-width increment (increment, keeping the word unchanged on overflow) \(i\) times from the all-false word of length \(w\) yields the width-\(w\) little-endian binary word of \(i\). Nothing is asserted at the terminal rank \(2^w\).
Writing a bit at position \(|\mathit{pre}|\) of the buffer tape holding \(\mathit{pre} \mathbin {+\! \! +} \mathit{old}\) yields the buffer tape holding \(\mathit{pre}\), then that bit, then the tail of \(\mathit{old}\) with its first cell dropped; this includes the case of an empty \(\mathit{old}\).
Let \(k \le K\) and suppose a host machine’s transitions on embedded states are the padded transitions of a \(k\)-tape source machine, reading only the first \(k\) tapes. Then for every configuration \(c\) of the source and every time \(t\), the host’s run from the padded configuration of \(c\) equals the padding of the source’s run.
For \(0 \lt k \le K\): padding the canonical configuration whose only nonblank tape is tape zero carrying a word \(s\) from a \(k\)-tape bank to a \(K\)-tape bank yields the canonical configuration with \(s\) on tape zero of the larger bank, all added tapes blank and parked at zero.
Given a map \(\iota \) from source control states to host control states, the padding of an action \(a\) of a \(k\)-tape machine to \(K\) tapes is the action that performs \(a\)’s input move, output, writes and moves on the first \(k\) work tapes, leaves every further tape unwritten and stationary, and replaces \(a\)’s next control state by its image under \(\iota \) (halting stays halting).
Consider a machine whose rewind state moves both endpoint heads left while the tape-zero cell is nonblank and exits right into an arrival state at the left blank. For every word \(w\) on tape zero and every \(j \le |w|\): from common head position \(j - 1\), exactly \(j + 1\) steps reach the arrival state with both heads at zero, preserving \(w\) and the blank capture tape.
A configuration on \(L + 1\) work tapes for the administrative copy and rewind scans: tape zero and the last tape carry given tape functions with a common head position, every other tape is blank with its head at zero, the input head is at cell one, and the output is empty.
Let \(M_1, M_2\) be finite machines, and let \(c_0\) be a configuration of \(M_1\) on input \(x\) (not necessarily initial) from which \(M_1\) halts within \(T_1\) steps with output \(y\); suppose \(M_2\) computes \(z\) from input \(y\) within \(T_2\) steps. Then the buffered composition of \(M_1\) and \(M_2\), started at its first-phase configuration for \(c_0\), halts within \(T_1 + |y| + 2 + T_2\) steps with output \(z\).
For any configuration \(c\), action \(a\), and work-tape index \(i\): writing the symbol that was under head \(i\) in \(c\), at \(c\)’s head position, into tape \(i\) after \(a\) has been applied restores tape \(i\) exactly to its contents in \(c\). This covers no-write actions and writes of the blank symbol.
Suppose a host machine’s transitions on embedded states are the halt-redirected transitions of a source machine, and the source is live at every time before \(t\) when run from \(c_0\). Then the host’s \(t\)-step run from the redirected configuration of \(c_0\) equals the redirected configuration of the source’s time-\(t\) configuration; in particular at the first halt the host lands in the return state with the source’s final writes and output.
Consider a machine with a start state whose transition unconditionally steps the input head left into a scanning state, and a scanning state that keeps stepping left over input symbols and exits right into a destination state at the left blank. Then from any configuration in the start state there is a number of steps, at most the current input position plus \(2\), after which the configuration equals the original one with control transferred to the destination state and the input head at cell one.
For any verifier machine \(MV\), words \(x, s\), and assembly phase: the buffered composition’s first-phase configuration over the assembly machine prepared with candidate \(s\) equals the canonical configuration in which control sits in the chosen assembly phase, tape zero carries \(s\), and all other tapes (composition buffer and verifier work) are blank.
The disjoint union of three finite machines on a common bank whose size is the sum of their tape counts: each branch runs its machine, padded, on an initial block of tapes. The machine starts in the third routine, and control can enter either other routine without disturbing tape zero.
The tape function determined by a finite list of possibly blank cells: a nonnegative position \(z\) holds the \(z\)-th list entry (which may itself be blank), and every position beyond the list, as well as every negative position, is blank.
For every \(p\), every bit \(b\), and every boolean word \(\mathit{rest}\), blanking position \(p\) of the sparse tape consisting of \(p\) blank cells followed by the symbols of the word \(b\) then \(\mathit{rest}\) yields the sparse tape of \(p + 1\) blank cells followed by the symbols of \(\mathit{rest}\).
Let \(c\) be a source configuration, \(a\) an action, \(h\) a history list, and \(p\) an input-head position. One step of the restoration machine from the checkpoint at \(p\) for the configuration obtained by applying \(a\) to \(c\), whose history is \(h\) extended by the entry recording \(a\) (the symbols under the work heads of \(c\) and the head moves of \(a\)), equals the mid-restoration configuration for \(c\), \(a\), \(h\) and \(p\): the recorded head moves are reversed, the overwritten symbols are retained in finite control, and the cells of that history entry are erased.
With no history left, one silent step of the restoration machine from the checkpoint for a source configuration \(c\) with empty history returns the history heads to zero and halts, reaching the completed restoration configuration for \(c\). This also covers a source run of zero steps.
Let \(M\) be a finite machine that is live (not halted) at every time before \(t\) when run from \(c_0\). Then the restoration machine, started at the checkpoint for \(M\)’s time-\(t\) configuration with the \(t\)-step history, reaches the completed restoration configuration of \(c_0\) — original work tapes and head positions, blank history — in exactly \(2t + 1\) steps, preserving any supplied input-head position.
Let \(c\) be a source configuration, \(a\) an action, \(h\) a history list, and \(p\) an input-head position. One step of the restoration machine from the mid-restoration configuration for \(c\), \(a\), \(h\) and \(p\) equals the checkpoint configuration for \(c\) and \(h\) at \(p\): the retained old symbols are written back, restoring the source tapes to their contents in \(c\), and the history heads move back onto the last entry of \(h\).
The checkpoint configuration of the restoration machine determined by a source configuration \(c\), a history list \(h\), and an input-head position: the source work fields are those of \(c\), the logged history tapes hold \(h\), and every history head inspects the last remaining entry.
The restoration entry action is the silent action that moves every history head of the restoration machine, including the clock head, one cell left, from the right blank onto the newest recorded entry; it writes nothing, leaves the source heads and the input head fixed, and selects the restoration machine’s checkpoint state.
The intermediate restoration configuration between inverse movement and inverse writing, for a source configuration \(c\), an action \(a\), and a history \(h\): the old symbols of \(c\) are retained in finite control, the source heads are back at their pre-action positions, and the last history entry is already erased.
For a source configuration \(c\) and an input-head position \(p\), the completed restoration configuration is the halted configuration of the restoration machine whose input head is at \(p\), whose source work tapes and head positions are those of \(c\), whose clock and history tapes are all blank with heads at zero, and whose output is empty.
The restoration controller on \(k + (1 + (k + k))\) work tapes: reading the newest history entry, it alternates moving the \(k\) source heads by the inverses of the recorded moves with writing back the recorded old symbols, erasing the history as it goes, and halts with all history heads at zero when the left blank of the clock tape is reached. The native input head never moves.
The inverse-move decoding is the map from recorded move symbols to head moves, used by the backward restoration pass, that sends a false bit to a right move, the blank symbol to staying put, and a true bit to a left move; thus every move symbol is decoded to the inverse of the move it encodes.
If a machine \(MV\) decides a language \(V\) within time \(T_v\), then from the canonical candidate seam with candidate \(s\) on input \(x\), the buffered composition of the assembly machine with \(MV\) halts within \(T_v(|x| + |s|) + 2(|x| + |s|) + 4\) steps, with single-bit output the indicator of \(x \mathbin {+\! \! +} s \in V\).
For every boolean word \(u\): if the fixed-width increment of \(u\) succeeds with result \(v\), then \(|v| = |u|\) and the little-endian value of \(v\) is one more than that of \(u\); if it overflows, then the value of \(u\) plus one equals \(2^{|u|}\).
Let \(M\) be a finite multi-tape Turing machine over the binary alphabet, \(x\) an input, and let \(\mathrm{cfg}\) assign to each rank a configuration of \(M\) on \(x\). Suppose the configuration at rank \(i + \mathrm{count}\) is halted with single-bit output false, and that for each rank \(j\) with \(i \le j \lt i + \mathrm{count}\) the machine, within \(B\) steps from \(\mathrm{cfg}(j)\), either halts with single-bit output true (when rank \(j\) accepts) or reaches \(\mathrm{cfg}(j+1)\) (when it rejects). Then within \(\mathrm{count} \cdot B\) steps from \(\mathrm{cfg}(i)\) the machine halts with single-bit output equal to the disjunction of the acceptance bits over the ranks \(i \le j \lt i + \mathrm{count}\).
Fix constants \(C, c\), a language \(V\) over the binary alphabet, and a machine \(MV\) deciding \(V\) within time \(T_v\). Then there exist a constant \(b\) and one finite machine \(E\) such that for every input \(x\) — writing \(n = |x|\) and \(w = C(n+1)^c\) — there is a rank-indexed family of configurations of \(E\) on \(x\) with: \(E\) reaches the rank-\(0\) configuration from its initial configuration within \(b\, (T_v(n+w) + (n+w+1)^{c+1})\) steps; the rank-\(2^w\) configuration is halted with single-bit output false; and from each rank \(i \lt 2^w\), within the same budget, \(E\) either halts with output true when \(x\) concatenated with the width-\(w\) binary word of \(i\) lies in \(V\), or reaches the rank-\((i+1)\) configuration otherwise.
Let \(\varphi \) be a CNF formula whose variables are indexed by natural numbers, and let \(N(\varphi )\) be one plus the largest variable index occurring in \(\varphi \) (\(0\) if no variable occurs). If two assignments \(a, b : \mathbb {N} \to \{ 0,1\} \) satisfy \(a(v) = b(v)\) for every \(v \lt N(\varphi )\), then \(\varphi \) evaluates to the same truth value under \(a\) as under \(b\).
For every \(\ell \in \mathbb {N}\) and every Boolean function \(f : \{ 0,1\} ^\ell \to \{ 0,1\} \) there exists a CNF formula \(\varphi \) over variables indexed by natural numbers such that: every variable occurring in \(\varphi \) has index less than \(\ell \) (one plus the largest occurring index, or \(0\) if none occurs, is at most \(\ell \)); \(\varphi \) has at most \(2^\ell \) clauses; every clause of \(\varphi \) contains at most \(\ell \) literals; and for every assignment \(a : \mathbb {N} \to \{ 0,1\} \), the formula \(\varphi \) evaluates under \(a\) to \(f(a(0), a(1), \dots , a(\ell -1))\) [AB09, Claim 2.13].
Let a machine \(MV\) decide a language \(V\) over the binary alphabet within time \(T_v\), and fix constants \(C, c\). Then there exist a constant \(b\) and a finite machine deciding the existential projection \(\{ \, x \mid \exists u,\ |u| = C(|x|+1)^c \wedge x \mathbin {+\! \! +} u \in V\, \} \) within time
A language \(L\) over the binary alphabet belongs to \(\mathrm{EXP}\) if and only if \(L \in \mathrm{DTIME}(2^{n^c})\) for some constant \(c \in \mathbb {N}\), i.e. \(L\) is decided by some finite multi-tape Turing machine within \(a \cdot 2^{n^c}\) steps on inputs of length \(n\), for constants \(a\) and \(c\).
Let \(T : \mathbb {N} \to \mathbb {N}\) be exponential-polynomial, i.e. \(T(n) \le 2^{K(n+1)^k}\) for all \(n\) and some constants \(K, k \in \mathbb {N}\). Then for all constants \(a, d \in \mathbb {N}\), the function \(n \mapsto T\bigl(a(n+1)^d\bigr)\) is also exponential-polynomial.
Let \(T : \mathbb {N} \to \mathbb {N}\) be exponential-polynomial, i.e. \(T(n) \le 2^{K(n+1)^k}\) for all \(n\) and some constants \(K, k \in \mathbb {N}\), and let \(L\) be a language over the binary alphabet in \(\mathrm{DTIME}(T)\), i.e. decided by some finite multi-tape Turing machine within \(c \cdot T(n)\) steps on inputs of length \(n\), for some constant \(c\). Then \(L \in \mathrm{EXP}\), that is, \(L \in \mathrm{DTIME}(2^{n^{c'}})\) for some constant \(c' \in \mathbb {N}\).
Let \(T_1, T_2 : \mathbb {N} \to \mathbb {N}\) both be exponential-polynomial, i.e. each is bounded by \(2^{K(n+1)^k}\) for all \(n\) and some constants \(K, k \in \mathbb {N}\). Then the product \(n \mapsto T_1(n) \cdot T_2(n)\) is exponential-polynomial.
Let \(T, T' : \mathbb {N} \to \mathbb {N}\) satisfy \(T(n) \le T'(n)\) for every \(n\). If \(T'\) is exponential-polynomial, i.e. \(T'(n) \le 2^{K(n+1)^k}\) for all \(n\) and some constants \(K, k \in \mathbb {N}\), then \(T\) is exponential-polynomial as well.
For all constants \(c, e \in \mathbb {N}\), the function \(n \mapsto c \cdot 2^{n^e}\) is exponential-polynomial: there exist constants \(K, k \in \mathbb {N}\) such that \(c \cdot 2^{n^e} \le 2^{K(n+1)^k}\) for every \(n \in \mathbb {N}\).
Let \(\ell \in \mathbb {N}\) and \(v = (v_0, \dots , v_{\ell -1}) \in \{ 0,1\} ^\ell \). The clause excluding \(v\) is
the disjunction, in order of \(i\), of \(\ell \) literals on the variables \(x_0, \dots , x_{\ell -1}\), where the literal \(x_i^{\neg v_i}\) is satisfied by an assignment \(a : \mathbb {N} \to \{ 0,1\} \) exactly when \(a(i) \ne v_i\) [AB09, Claim 2.13].
Let \(\ell \in \mathbb {N}\), let \(v \in \{ 0,1\} ^\ell \), and let \(C_v = \bigvee _{i\lt \ell } x_i^{\neg v_i}\) be the clause on the variables \(x_0, \dots , x_{\ell -1}\) whose \(i\)-th literal is satisfied by an assignment \(a\) exactly when \(a(i) \ne v_i\). For every assignment \(a : \mathbb {N} \to \{ 0,1\} \), the clause \(C_v\) evaluates to false under \(a\) if and only if \(a(i) = v_i\) for all \(i \lt \ell \) [AB09, Claim 2.13].
Let \(\ell \in \mathbb {N}\) and let \(f : \{ 0,1\} ^\ell \to \{ 0,1\} \). The truth-table CNF \(\varphi _f\) of \(f\) is the conjunction, over all \(v \in \{ 0,1\} ^\ell \) with \(f(v) = 0\), of the clause \(C_v = \bigvee _{i\lt \ell } x_i^{\neg v_i}\), whose \(i\)-th literal is satisfied by an assignment \(a : \mathbb {N} \to \{ 0,1\} \) exactly when \(a(i) \ne v_i\) [AB09, Claim 2.13].
Let \(\ell \in \mathbb {N}\), let \(f : \{ 0,1\} ^\ell \to \{ 0,1\} \), and let \(\varphi _f\) be the CNF formula that is the conjunction, over all \(v \in \{ 0,1\} ^\ell \) with \(f(v) = 0\), of the clause \(\bigvee _{i\lt \ell } x_i^{\neg v_i}\) whose \(i\)-th literal is satisfied by an assignment \(a\) exactly when \(a(i) \ne v_i\). Then for every assignment \(a : \mathbb {N} \to \{ 0,1\} \), the formula \(\varphi _f\) evaluates under \(a\) to \(f(a(0), a(1), \dots , a(\ell -1))\).
Let \(\ell \in \mathbb {N}\), let \(f : \{ 0,1\} ^\ell \to \{ 0,1\} \), and let \(\varphi _f\) be the CNF formula that is the conjunction, over all \(v \in \{ 0,1\} ^\ell \) with \(f(v) = 0\), of the clause \(\bigvee _{i\lt \ell } x_i^{\neg v_i}\) whose \(i\)-th literal is satisfied by an assignment \(a\) exactly when \(a(i) \ne v_i\). Then \(\varphi _f\) has at most \(2^\ell \) clauses.
Let \(\ell \in \mathbb {N}\), let \(f : \{ 0,1\} ^\ell \to \{ 0,1\} \), and let \(\varphi _f\) be the CNF formula that is the conjunction, over all \(v \in \{ 0,1\} ^\ell \) with \(f(v) = 0\), of the clause \(\bigvee _{i\lt \ell } x_i^{\neg v_i}\) whose \(i\)-th literal is satisfied by an assignment \(a\) exactly when \(a(i) \ne v_i\). Then every variable occurring in \(\varphi _f\) has index less than \(\ell \); that is, one plus the largest variable index occurring in \(\varphi _f\) (or \(0\) if none occurs) is at most \(\ell \).
Let \(\ell \in \mathbb {N}\), let \(f : \{ 0,1\} ^\ell \to \{ 0,1\} \), and let \(\varphi _f\) be the CNF formula that is the conjunction, over all \(v \in \{ 0,1\} ^\ell \) with \(f(v) = 0\), of the clause \(\bigvee _{i\lt \ell } x_i^{\neg v_i}\) whose \(i\)-th literal is satisfied by an assignment \(a\) exactly when \(a(i) \ne v_i\). Then every clause of \(\varphi _f\) contains at most \(\ell \) literals, i.e. \(\varphi _f\) is an \(\ell \)CNF.
Let \(\alpha = \alpha _1 \cdots \alpha _m\) be a binary string, and for a binary string \(x\) let \(\langle \alpha , x \rangle = \alpha _1\alpha _1 \cdots \alpha _m\alpha _m\, 0\, 1\, x\) be the pairing encoding (\(\alpha \) with every bit doubled, then the separator \(01\), then \(x\)). The machine with no work tapes that first emits the fixed word \(\alpha _1\alpha _1 \cdots \alpha _m\alpha _m\, 01\) one bit per step and then copies its input bit by bit, halting at the first blank, halts on every input \(x\) within \(2|\alpha | + |x| + 3\) steps with output \(\langle \alpha , x \rangle \).
Let \(\alpha \) be a fixed binary string, and for binary strings \(x\) let \(\langle \alpha , x \rangle \) be the pairing encoding: \(\alpha \) with every bit doubled, followed by \(01\), followed by \(x\). Then the map \(x \mapsto \langle \alpha , x \rangle \) is polynomial-time computable: some finite-state binary-alphabet Turing machine computes it within \(C \cdot (n+1)^c\) steps on inputs of length \(n\), for constants \(C, c \in \mathbb {N}\).
Let \(s\) be a finite list of natural numbers, write \(\max (s)\) for the right fold of \(\max \) over \(s\) with initial value \(0\), and let \(n \in \mathbb {N}\). If every member of \(s\) is at most \(n\), then \(\max (s) \le n\); in particular this holds when \(s\) is empty.
Given a direction (rightward or leftward), a guide tape \(\gamma \), and \(j \in \mathbb {N}\), the \(j\)-cell filled version of \(\gamma \) is the tape agreeing with \(\gamma \) except that each cell \(z\) with \(0 \le z \lt j\) (rightward) respectively \(0 \le -z \lt j\) (leftward) is overwritten with the guide mark of \(z\): the origin mark at \(0\) and interior marks elsewhere, independently of the previous contents.
Let \(R \gt 0\). Starting from the right guide of radius \(R\), filling \(R\) cells leftward and then writing an interior mark at cell \(-R\) and the left outer boundary marker at cell \(-R-1\) produces the completed guide of radius \(R\): origin mark at \(0\), interior marks at all \(z\) with \(0 \lt |z| \le R\), boundary markers at \(\pm (R+1)\), and blanks elsewhere.
For a radius \(R\), the guide tape carries a left edge marker at coordinate \(-R - 1\), a right edge marker at coordinate \(R + 1\), a marked origin at coordinate \(0\), an interior marker at every other coordinate \(z\) with \(-R \le z \le R\), and blanks elsewhere.
Fix a coding scheme \(c\) that assigns to every bit string \(\alpha \) a Turing machine (the machine that \(\alpha \) denotes under \(c\)). The halting function \(\mathrm{HALT}_c\) maps a bit string \(s\) to a Boolean: \(\mathrm{HALT}_c(s)\) is true precisely when \(s\) is the pairing encoding \(\langle \alpha , x \rangle \) of two bit strings — the universal machine’s input format, code first — and the machine denoted by \(\alpha \) halts on input \(x\), i.e. completes some output within some finite number of steps. Strings not of this form (the pairing encoding is not surjective) are mapped to false.
Let \(c\) be a coding scheme for Turing machines and let \(s\) be a bit string. Then \(\mathrm{HALT}_c(s)\) is true if and only if there exist bit strings \(\alpha \) and \(x\) such that \(s\) is the pairing encoding \(\langle \alpha , x \rangle \) and the machine that \(\alpha \) denotes under \(c\) produces some output on input \(x\) within some finite number of steps.
Let \(c\) be an effective coding scheme for Turing machines, i.e. one admitting a universal evaluator machine, and let \(\mathrm{HALT}_c\) be its halting function: true on a bit string \(s\) exactly when \(s\) is the pairing encoding \(\langle \alpha , x \rangle \) of two bit strings and the machine denoted by \(\alpha \) under \(c\) produces some output on input \(x\) within some finite number of steps, and false otherwise. Then no Turing machine computes the function mapping \(s\) to the one-bit string carrying \(\mathrm{HALT}_c(s)\) [AB09, Theorem 1.11].
Let \(c\) be an effective representation scheme for Turing machines: a total decoding of bit strings into coded machines with an encoding that decoding inverts (also after padding the code with \(1\)s), equipped with a finite-state binary-alphabet machine that, within some time bound, computes from every bit string the canonical serialization of the machine it decodes to. Let \(\mathrm{HALT}_c\) be the set of bit strings of the form \(\langle \alpha , x \rangle \) (\(\alpha \) with every bit doubled, then \(01\), then \(x\)) such that the machine decoded from \(\alpha \) halts on input \(x\). Then \(\mathrm{HALT}_c \notin \mathsf{NP}\) [AB09, Exercise 2.8].
Let \(c\) be any representation scheme for Turing machines: a total decoding of bit strings into coded one-work-tape machines together with an encoding, such that a machine’s code followed by any number of \(1\)s decodes back to that machine. Let \(\mathrm{HALT}_c\) be the set of bit strings of the form \(\langle \alpha , x \rangle \) (\(\alpha \) with every bit doubled, then \(01\), then \(x\)) such that the machine decoded from \(\alpha \) halts on input \(x\). Then \(\mathrm{HALT}_c\) is \(\mathsf{NP}\)-hard: every language in \(\mathsf{NP}\) is polynomial-time Karp reducible to \(\mathrm{HALT}_c\) [AB09, Exercise 2.8].
Let \(c\) be a coding scheme for Turing machines and let \(\alpha \) and \(x\) be bit strings. Then \(\mathrm{HALT}_c\), evaluated on the pairing encoding \(\langle \alpha , x \rangle \), is true if and only if the machine that \(\alpha \) denotes under \(c\) halts on input \(x\), completing some output within some finite number of steps.
For a binary machine \(M\), a payload assignment \(f\), a finite data register, data tapes, and an output list, the initial-content predicate states that the register is the initial data state of \(M\), the output is empty, and the current payload decoded from the data tapes at every lane and coordinate equals the value of \(f\) there.
Let \(P\) be a machine over an alphabet \(A\) and let \(z \in A\). Suppose the transition table of \(P\) is insensitive to nonblank input contents: for every state \(q\), scanned input symbol \(s\), and tuple of work-tape reads \(w\), the transition of \(P\) on the read obtained from \(s\) by replacing any nonblank symbol with \(z\) coincides with the transition of \(P\) on \(s\) itself. Then \(P\) is oblivious.
For a binary word \(x\) and an integer coordinate \(z\), the virtual input payload at \(z\) consists of the buffered input symbol of \(x\) at \(z\) together with the tag \(1\) at \(z = -1\) (the left boundary), the tag \(2\) at \(z = |x|\) (the right boundary), and the ordinary tag \(0\) elsewhere.
For every configuration of a binary machine on an input \(x\) with input head at position \(p\), the virtual input payload of \(x\) at coordinate \(p - 1\) is exactly the pair of the symbol read by the native input head and the boundary tag of that configuration; on the empty input both boundary tags are produced correctly.
Let \(M\) be a binary multi-tape Turing machine and let \(n, t \in \mathbb {N}\). On an input of length \(n\) the input cells are numbered \(0, 1, \dots , n+1\), where cell \(p\) with \(1 \le p \le n\) holds the \(p\)-th input bit and cells \(0\) and \(n+1\) are blank boundary cells. The scheduled input position \(\mathrm{inpos}_M(n, t) \in \{ 0, \dots , n+1\} \) is the position of the input head after \(t\) steps of \(M\) on the all-zeros input \(0^n\).
For a configuration of a binary machine on an input \(x\), the boundary tag at the native input head is \(1\) when the head sits at position \(0\) (the left end), \(2\) when it sits at position \(|x| + 1\) (the right end), and \(0\) otherwise.
Let \(L_1, L_2 \subseteq \{ 0,1\} ^*\) be languages in \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time. Then their intersection \(L_1 \cap L_2 = \{ z : z \in L_1 \text{ and } z \in L_2\} \) also belongs to \(\mathsf{P}\).
Let a \(k\)-tape machine, a designated lane, an encoding of transducer states into machine states, an encoding of transducer letters into tape symbols, and a local transducer step be given. Suppose that from every encoded transducer state, whenever the lane reads an encoded letter, the machine performs exactly the single-lane action that writes the encoded output letter of the local step, moves right, and enters the encoded updated transducer state. Then running the machine for exactly \(\ell \) steps from a single-lane configuration whose unprocessed list begins with an encoded word of length \(\ell \) yields the single-lane configuration in which the transducer has been folded over the whole word: the state is the encoded final transducer state, the coordinate has advanced by \(\ell \), the encoded output word has been moved in reverse onto the processed list, and no inactive tape has changed.
Given a designated work tape, a control state, an optional symbol, and a direction, the single-lane action writes that symbol and moves in that direction on the designated tape only, enters the given state, keeps the input head stationary, leaves every other work tape untouched, and emits no output.
Let a base configuration of a \(k\)-tape machine be given, together with a distinguished work tape (the lane), an optional control state, an integer coordinate \(z\), and two lists of optional symbols: the already-processed cells in reverse order and the unprocessed cells. The lane configuration agrees with the base on every other tape, on the input head, and on the output, while the lane carries the zipper tape determined by \(z\) and the two lists and its head sits at \(z\).
Let a single-lane configuration be given whose unprocessed list begins with an entry \(s\). Applying the single-lane action that writes an entry \(s'\) on the lane, moves right, and enters a state \(q\) yields again a single-lane configuration: the state becomes \(q\), the coordinate advances by one, \(s'\) is pushed onto the processed list, \(s\) is removed from the unprocessed list, and all other tapes and heads are unchanged.
Take a setup configuration and focus its budget lane: replace the state by \(q'\), the budget tape by the zipper tape with frontier \(z\), left list \(l\), and right list \(r\), and move the budget head to \(z\), keeping every other tape and head. The result is again a setup configuration, with the same base configuration, input head, unary and guide data, state \(q'\), budget head \(z\), and budget tape given by that zipper.
For every bit list \(w\) and every optional bit \(b\), the last entry, defaulting to false, of \(w\) extended by the zero- or one-element list determined by \(b\) equals \(b\) when \(b\) is present, and the last entry of \(w\) defaulting to false otherwise; a one-bit last-output register can therefore be maintained in finite control.
Let \(\pi _1\) and \(\pi _2\) be the total projections on bit strings, which send a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(a\) and to \(b\) respectively, and send every other bit string to the empty word. Then the language \(\{ z \in \{ 0,1\} ^* : |\pi _1(z)| = |\pi _2(z)|\} \) belongs to \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time.
Let \(f, g : \{ 0,1\} ^* \to \{ 0,1\} ^*\) be polynomial-time computable functions (each computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then the language \(\{ z \in \{ 0,1\} ^* : |f(z)| = |g(z)|\} \) belongs to \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time.
Let \(\pi _1\) be the total first projection, which sends a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(a\), and any other bit string to the empty word. Then for every bit string \(z\), \(|\pi _1(z)| \le |z|\).
Let \(\pi _1\) and \(\pi _2\) be the total projections on bit strings, which send a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(a\) and to \(b\) respectively, and send every other bit string to the empty word. Then the language \(\{ z \in \{ 0,1\} ^* : |\pi _2(z)| \le |\pi _1(z)|\} \) belongs to \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time.
Let \(f, g : \{ 0,1\} ^* \to \{ 0,1\} ^*\) be polynomial-time computable functions (each computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then the language \(\{ z \in \{ 0,1\} ^* : |g(z)| \le |f(z)|\} \) belongs to \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time.
Let \(W\) be a binary machine, \(a\) a padding constant, and let a base configuration of the schedule controller be given together with an optional control phase and integers \(u\) and \(g\). The macrostep configuration agrees with the base on every tape content, on the input head, on the budget head, and on the output; only the control phase (set to the given one), the unary-counter head (at \(u\)), and the guide head (at \(g\)) vary.
Let a macrostep configuration with phase \(q\), unary head \(u\), and guide head \(g\) be given. Applying the read-only schedule action that enters a phase \(q'\), keeps the input and budget heads fixed, writes nothing, and moves the unary and guide heads in directions \(d_u\) and \(d_g\) yields the macrostep configuration with phase \(q'\), unary head \(u + d_u\), and guide head \(g + d_g\).
Suppose the base’s guide tape is the fixed guide of radius \(R\). From the backward phase with guide head at \(g\), one step of the schedule controller turns into the return-to-center phase and moves the guide head right exactly when \(g = -R - 1\), and otherwise stays in the backward phase and moves the guide head left.
Suppose the base’s guide tape is the fixed guide of radius \(R\) and let \(n \le 2R + 1\). Running the schedule controller for \(n\) steps from the backward phase with guide head at \(R\) stays in the backward phase and moves the guide head to \(R - n\).
Suppose the base’s guide tape is the fixed guide of radius \(R\). From the return-to-center phase with guide head at \(g\), one step of the schedule controller commits exactly at the origin: when \(g = 0\) it enters the counter-test phase and advances the unary head by one, and otherwise it stays in the return phase and moves the guide head right.
Suppose the base’s guide tape is the fixed guide of radius \(R\) and let \(n \le R\). Running the schedule controller for \(n\) steps from the return-to-center phase with guide head at \(-R\) stays in the return phase and moves the guide head to \(-R + n\).
For every macrostep configuration in the counter-test phase with unary head \(u\) and guide head \(g\): if the base’s unary tape holds a unit marker at \(u\), one step of the schedule controller enters the leftward-seek phase; otherwise it halts. In both cases every head stays in place.
Suppose the base’s guide tape is the fixed guide of radius \(R\) and its unary tape holds a unit marker at coordinate \(u\). Then running the schedule controller for exactly \(6R + 8\) steps from the counter-test phase with unary head at \(u\) and guide head at the origin executes the entire fixed path of one macrostep and returns to the counter-test phase with unary head at \(u + 1\) and guide head at the origin.
Suppose the base’s guide tape is the fixed guide of radius \(R\), its unary tape holds unit markers at all coordinates \(1, 2, \dots , B\), and the unary cell at coordinate \(B + 1\) is blank. Then running the schedule controller for \(B\, (6R + 8) + 1\) steps from the counter-test phase with unary head at \(1\) and guide head at the origin ends in the halted phase with unary head at \(B + 1\) and guide head at the origin: the schedule finishes after exactly the budgeted macrosteps plus one final counter test.
Suppose the base’s guide tape is the fixed guide of radius \(R\). From the forward phase with guide head at \(g\), one step of the schedule controller turns into the backward phase and moves the guide head left exactly when \(g = R + 1\), and otherwise stays in the forward phase and moves the guide head right.
Suppose the base’s guide tape is the fixed guide of radius \(R\) and let \(n \le 2R + 1\). Running the schedule controller for \(n\) steps from the forward phase with guide head at \(-R\) stays in the forward phase and moves the guide head to \(-R + n\).
Suppose the base’s guide tape is the fixed guide of radius \(R\) and its unary tape holds unit markers at all coordinates \(1, 2, \dots , B\). Then for every \(j \le B\), running the schedule controller for \(j\, (6R + 8)\) steps from the counter-test phase with unary head at \(1\) and guide head at the origin reaches the counter-test phase with unary head at \(j + 1\) and guide head at the origin, regardless of all simulated data and of source halting behavior.
Suppose the base’s guide tape is the fixed guide of radius \(R\). From the leftward-seek phase with guide head at \(g\), one step of the schedule controller turns into the forward phase and moves the guide head right exactly when \(g = -R - 1\), and otherwise stays in the seek phase and moves the guide head left.
Suppose the base’s guide tape is the fixed guide of radius \(R\) and let \(n \le R + 1\). Running the schedule controller for \(n\) steps from the leftward-seek phase with guide head at the origin stays in the seek phase and moves the guide head to \(-n\).
Modifying a setup configuration by changing only its phase to \(q'\) and its unary and guide head positions to \(u'\) and \(g'\) — keeping every tape, the budget head, and the input head fixed — again yields a setup configuration with the same remaining data.
Given a symbol \(z\) and a configuration \(c\) of a \(k\)-tape machine on an input word \(x\), the masked configuration is the configuration on the constant input \(z^{\left\lvert x\right\rvert }\) (the word consisting of \(\left\lvert x\right\rvert \) copies of \(z\)) that has the same control state, the same numerical input-head position, the same work-tape contents and head positions, and the same output stream as \(c\).
Let \(z\) be a symbol, \(a\) an action (an input-head move, a write and a move for each work tape, an optional emitted symbol, and a next state), and \(c\) a configuration on an input word \(x\). Then masking the input by \(z\) after applying \(a\) yields the same configuration as applying \(a\) to the masked configuration.
Let \(z\) be a symbol and \(c\) a configuration of a \(k\)-tape machine on an input word \(x\). The symbol scanned by the input head of the masked configuration on \(z^{\left\lvert x\right\rvert }\) is blank exactly when the symbol scanned in \(c\) is blank, and equals \(z\) whenever \(c\) scans any nonblank symbol.
Given a machine \(W\) over an alphabet \(A\) and a distinguished symbol \(z\), the masked clock is the machine with the same number of tapes, the same state set, and the same start state as \(W\), whose transition function first replaces every nonblank scanned input symbol by \(z\) (leaving blank reads blank) and then consults \(W\)’s transition table. Its behaviour on any input therefore matches \(W\)’s behaviour on the constant input of the same length.
Let \(W\) be a machine over the Boolean alphabet, \(T : \mathbb {N} \to \mathbb {N}\), and \(b \in \mathbb {N}\). Suppose that on every input \(x\), after \(b \cdot (T(\left\lvert x\right\rvert ) + 1)\) steps \(W\) has halted with output exactly the list of binary digits of \(T(\left\lvert x\right\rvert )\). Then the masked clock of \(W\) (with mask the symbol \(\mathsf{false}\)) satisfies the same specification: on every input \(x\) it halts within \(b \cdot (T(\left\lvert x\right\rvert ) + 1)\) steps with output the binary digits of \(T(\left\lvert x\right\rvert )\).
Let \(W\) be a machine over an alphabet \(A\), let \(z \in A\), let \(x\) be an input word, and let \(t \in \mathbb {N}\). Masking the configuration reached by the masked clock of \(W\) after \(t\) steps from its initial configuration on \(x\) yields exactly the configuration reached by \(W\) after \(t\) steps from its initial configuration on the constant input \(z^{\left\lvert x\right\rvert }\) — an equality of all five configuration components (state, input position, work tapes, work-head positions, and output), at live and halted times alike.
For every machine \(W\) over an alphabet \(A\) and every symbol \(z \in A\), the masked clock of \(W\) with mask \(z\) is oblivious: on any two inputs of the same length its input-head position and all of its work-head positions agree at every time. No halting hypothesis is required.
Let \(W\) be a machine over an alphabet \(A\), let \(z \in A\), and let \(c\) be a configuration of the masked clock of \(W\) on an input word \(x\). Masking the result of one step of the masked clock at \(c\) yields the same configuration as performing one step of \(W\) at the masked configuration of \(c\); this includes the absorbing case in which \(c\) is already halted.
A language \(L \subseteq \{ 0,1\} ^{*}\) belongs to the class coNP (that is, its complement lies in \(\mathsf{NP}\)) if and only if there exist natural numbers \(C\) and \(c\) and a verifier language \(V \in \mathsf{P}\) such that for every binary string \(x\),
where \(xu\) denotes the concatenation of \(x\) and \(u\).
Let \(\Gamma \) be a finite alphabet, let \(e : \{ 0,1\} \hookrightarrow \Gamma \) be an injective embedding of the binary symbols, let \(L \subseteq \{ 0,1\} ^{*}\) be a language over the binary alphabet, and let \(T : \mathbb {N} \to \mathbb {N}\) be a time bound. If some deterministic finite Turing machine over \(\Gamma \) decides \(L\) via \(e\) within time \(T\), then \(L \in \mathrm{DTIME}(n \mapsto T(n) + 1)\); that is, some deterministic finite Turing machine over the binary alphabet decides \(L\) within \(a \cdot (T(n) + 1)\) steps on every input of length \(n\), for some constant \(a\).
For a language \(L \subseteq \{ 0,1\} ^{*}\), the following are equivalent: (i) \(L \in \mathsf{NP}\), i.e. there are \(C', c' \in \mathbb {N}\) and a language \(V' \in \mathsf{P}\) such that for every binary string \(x\), \(x \in L\) if and only if some \(u\) with \(|u| = C'(|x|+1)^{c'}\) has \(xu \in V'\); (ii) there are \(C, c \in \mathbb {N}\) and a language \(V \in \mathsf{P}\) such that for every binary string \(x\), \(x \in L\) if and only if some binary string \(u\) with \(|u| \le C(|x|+1)^c\) has \(\langle x, u\rangle \in V\), where \(\langle x, u\rangle \) is the self-delimiting pair encoding consisting of \(x\) with every bit doubled, then the separator \(01\), then \(u\) verbatim.
A language \(L\) over the binary alphabet belongs to \(\mathsf{P}\) if and only if there exist natural numbers \(C\) and \(d\) and a deterministic finite Turing machine \(M\) over \(\{ 0,1\} \) such that \(M\) decides \(L\) within \(C \cdot (n + 1)^d\) steps on every input of length \(n\). In particular the \(+1\) padding in the definition of \(\mathsf{P}\) describes the same class as the conventional polynomial time bounds \(C \cdot (n+1)^d\).
A language \(L\) over the binary alphabet belongs to \(\mathsf{P}\) if and only if there exist a deterministic finite Turing machine \(M\) over \(\{ 0,1\} \) and natural numbers \(C\) and \(d\) such that \(M\) has exactly one work tape and \(M\) decides \(L\) within \(C \cdot (n+1)^d\) steps on every input of length \(n\). Thus restricting to a single work tape does not change the class \(\mathsf{P}\): the quadratic overhead of tape reduction preserves polynomial time bounds.
Let \(V \subseteq \{ 0,1\} ^*\) be a language and let \(\mathbf{1}_V(x) \in \{ 0,1\} \) be its indicator bit, equal to \(1\) if \(x \in V\) and \(0\) otherwise. Then \(V\) belongs to \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time, if and only if the map sending a bit string \(x\) to the one-bit word \(\mathbf{1}_V(x)\) is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
Let \(\iota \) be a finite index set and, for each \(i \in \iota \), let \(b_i : \{ 0,1\} ^* \to \{ 0,1\} \) be a Boolean test on bit strings such that the language \(\{ z : b_i(z) = 1\} \) belongs to \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time. Then for every Boolean function \(F : \{ 0,1\} ^{\iota } \to \{ 0,1\} \), the language
also belongs to \(\mathsf{P}\).
Let \(\Gamma \) be a finite alphabet, let \(e : \{ 0,1\} \hookrightarrow \Gamma \) be an injective embedding of the binary symbols, let \(L \subseteq \{ 0,1\} ^{*}\) be a language over the binary alphabet, and let \(C\) and \(d\) be natural numbers. If some deterministic finite Turing machine over \(\Gamma \) decides \(L\) via \(e\) within \(C \cdot (n+1)^d\) steps on every input of length \(n\), then \(L \in \mathsf{P}\). In particular, polynomial-time decidability over any finite alphabet already places a language in the binary-alphabet class \(\mathsf{P}\).
Let \(L\) be a language over the binary alphabet, let \(T : \mathbb {N} \to \mathbb {N}\) be a time bound with \(L \in \mathrm{DTIME}(T)\), and let \(c\) and \(d\) be natural numbers such that \(T(n) \le c \cdot (n^d + 1)\) for every \(n\). Then \(L \in \mathsf{P}\). (The domination is required pointwise, at every input length, not merely for all sufficiently large \(n\).)
Let \(b : \{ 0,1\} ^* \to \{ 0,1\} \) be a Boolean test on bit strings such that the map sending \(z\) to the one-bit word \(b(z)\) is polynomial-time computable (computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then the language \(\{ z \in \{ 0,1\} ^* : b(z) = 1\} \) belongs to \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time.
A language \(L\) over the binary alphabet belongs to \(\mathrm{NEXP}\) if there exist constants \(C, c \in \mathbb {N}\) and a verifier language \(V \in \mathrm{P}\) such that for every binary string \(x\): \(x \in L\) if and only if there is a certificate \(u\) of length exactly \(C \cdot 2^{(|x|+1)^c}\) with the concatenation \(x \mathbin {+\! \! +} u\) in \(V\).
A language \(L \subseteq \{ 0,1\} ^{*}\) belongs to the class \(\mathsf{NP}\) if there exist natural numbers \(C\) and \(c\) and a verifier language \(V \in \mathsf{P}\) such that for every binary string \(x\), \(x \in L\) holds if and only if there is a certificate \(u \in \{ 0,1\} ^{*}\) of length exactly \(|u| = C \cdot (|x|+1)^c\) whose concatenation \(xu\) with \(x\) lies in \(V\).
If the class \(\mathsf{P}\) of polynomial-time decidable binary languages equals \(\mathsf{NP}\), then \(\mathsf{NP}\) equals coNP, where coNP is the class of languages \(L \subseteq \{ 0,1\} ^{*}\) whose complement \(\{ 0,1\} ^{*} \setminus L\) is in \(\mathsf{NP}\).
Every language in \(\mathrm{NP}\) lies in \(\mathrm{EXP}\): a language over the binary alphabet with polynomial-length certificates and a polynomial-time verifier is decidable in time \(2^{n^e}\) for some constant \(e\), up to a constant factor.
Let \(L \subseteq \{ 0,1\} ^{*}\) be an \(\mathsf{NP}\)-complete language, i.e. \(L \in \mathsf{NP}\) and every language in \(\mathsf{NP}\) is polynomial-time Karp reducible to \(L\). Then \(L \in \mathsf{P}\) if and only if \(\mathsf{P} = \mathsf{NP}\) [AB09, Theorem 2.8.3].
Let \(L\) and \(L'\) be languages over the Boolean alphabet. If every language in \(\mathrm{NP}\) polynomial-time many-one reduces to \(L\), and \(L\) polynomial-time reduces to \(L'\), then every language in \(\mathrm{NP}\) polynomial-time reduces to \(L'\).
For a time bound \(T : \mathbb {N} \to \mathbb {N}\), the class \(\mathrm{NTIME}(T)\) consists of all languages \(L\) of binary strings for which there exist a constant \(c \in \mathbb {N}\) and a finite nondeterministic multi-tape Turing machine \(N\) over the binary alphabet such that, on every input \(x\), every choice word of length \(c \cdot T(|x|)\) drives \(N\) to a halted configuration, and \(x \in L\) exactly when some choice word of that length drives \(N\) to a halted configuration with output \(1\).
Let \(T_1, T_2 : \mathbb {N} \to \mathbb {N}\) be time bounds with \(T_1(n) \le T_2(n)\) for every \(n\). Then \(\mathrm{NTIME}(T_1) \subseteq \mathrm{NTIME}(T_2)\): every language of binary strings decided by some finite nondeterministic multi-tape machine within \(c \cdot T_1(n)\) steps on inputs of length \(n\), for some constant \(c\), is also decided by such a machine within \(c \cdot T_2(n)\) steps.
Let \(T : \mathbb {N} \to \mathbb {N}\) be a time bound with \(T(n) = 0\) for at least one input length \(n\). Then \(\mathrm{NTIME}(T) = \emptyset \): no language of binary strings is decided by a finite nondeterministic multi-tape machine within \(c \cdot T(|x|)\) steps on every input \(x\), for any constant \(c\).
For every natural number \(c\), the class \(\mathrm{NTIME}(2^{n^c})\) is contained in the certificate-form class \(\mathrm{NEXP}\): every language decided by some nondeterministic Turing machine within \(a \cdot 2^{n^c}\) steps admits a polynomial-time verifier with certificates of exact length \(a \cdot 2^{(n+1)^c}\).
For every natural number \(c\), the nondeterministic time class \(\mathrm{NTIME}(n^c + 1)\) is contained in \(\mathrm{NP}\): every language decided by some nondeterministic Turing machine within time \(a\, (n^c+1)\), for a constant \(a\), admits a polynomial-time verifier with certificates of exact length \(2a\, (n+1)^c\).
An action builder for a machine with \(k + 3\) work tapes over the schedule alphabet. Given an optional next state (blank meaning halt), an input-head move, and a write–move pair for each of three extra tapes — the binary budget, the unary macrostep counter, and the sweep guide — it produces the action that leaves all \(k\) clock work tapes unwritten and stationary, performs the three given write–move pairs on the extra tapes, and emits no output.
Consider a machine with \(k\) leading work tapes followed by three administrative tapes, of which the first is the budget tape. For every state \(q\), optional symbol \(s\), and direction \(d\), the schedule action that enters \(q\), keeps the input head fixed, writes \(s\) with direction \(d\) on the budget tape, and leaves the two remaining administrative tapes and all \(k\) leading tapes untouched equals the single-lane action on the budget tape writing \(s\), moving in direction \(d\), and entering \(q\).
Let \(M\) be a binary multi-tape Turing machine with \(k\) work tapes. The associated finite data state consists of an optional control state of \(M\) (absent once \(M\) has halted), a saved one-bit answer, and two vectors of \(k+1\) payloads: the reads captured at macrostep entry and the neighbor registers used during a sweep. All tape coordinates remain on tapes, so this state is finite.
Let \(T\) be a time-constructible bound and let \(L\) be a binary language in \(\mathrm{DTIME}(T)\). Then there exist an oblivious machine \(M\) — one whose head positions at every time step depend only on the input length — and a constant \(c\) such that \(M\) decides \(L\) within time \(c\, (T(n)+1)^2\). This is the first assertion of Arora–Barak, Exercise 1.5.
Let \(M\) be a binary multi-tape Turing machine with \(k\) work tapes that is oblivious: for any two inputs of the same length and every \(t \in \mathbb {N}\), the input-head position and all work-head positions after \(t\) steps agree on the two inputs. Then for every input \(x \in \{ 0,1\} ^*\) and every \(t \in \mathbb {N}\), the input-head position after \(t\) steps of \(M\) on \(x\) equals the input-head position after \(t\) steps of \(M\) on the all-zeros input \(0^{|x|}\), and for every work tape \(\tau \) the head position of tape \(\tau \) after \(t\) steps on \(x\) equals its head position after \(t\) steps on \(0^{|x|}\).
Given binary machines \(W\) (the clock witness) and \(M\) (the source) and a padding constant \(a\), the candidate simulator is obtained by decorating the length-only schedule controller built from \(W\) and \(a\) with \(k+1\) additional data tapes driven by the data transducer, starting from the initial data state, and then realizing the finite alphabet over the binary alphabet by simultaneous transverse binary blocks; this coding retains the exact schedule.
Let \(W\) compute on every input \(x\) the binary expansion of \(T(|x|)\) within \(b\, (T(|x|)+1)\) steps, suppose \(n \le T(n)\) for all \(n\), and let \(M\) decide the language \(L\) within time \(a\, T(n)\). Then the candidate simulator built from \(W\), \(M\), and \(a\) decides \(L\) within time \(\bigl(18(a+1)^2 + 23(a+1) + 3b + 25\bigr)\, (T(n)+1)^2\).
Let \(W\) and \(M\) be machines over the Boolean alphabet, let \(a, b \in \mathbb {N}\), and let \(T \colon \mathbb {N}\to \mathbb {N}\) satisfy \(n \le T(n)\) for every \(n\). Suppose that on every Boolean input \(x\) the machine \(W\) halts within \(b\, (T(\left\lvert x\right\rvert )+1)\) steps with output exactly the list of (little-endian) binary digits of \(T(\left\lvert x\right\rvert )\). Then for every Boolean input \(x\), the candidate oblivious simulator built from \(W\), \(M\), and \(a\) — the length-only schedule decorated by the data transducer for \(M\) and realized over the binary alphabet by transverse binary blocks — halts on \(x\) within the same uniform quadratic budget: there exists a time
after which its control state is the halting state. No hypothesis on the computation of the source machine \(M\) is needed, and the bound does not depend on \(M\).
For all binary machines \(W\) and \(M\) and every padding constant \(a\), the candidate simulator built from \(W\), \(M\), and \(a\) is oblivious: its head positions at every time step depend only on the length of the input. This holds on all binary inputs of every length, including the empty input, and uses no assumption about the computations of \(W\) or \(M\).
Let \(a, b, u, n, \tau , w \in \mathbb {N}\) satisfy \(n \le u\), \(\tau \le b\, (u+1)\), and \(w \le b\, (u+1)\). Then
The left-hand side gathers, as abstract summands, the costs of a clock run of length \(\tau \), an input reset over \(n\) cells, the fixed overheads, \(u+1\) budget rounds over a word of width \(w\), and the allocation and sweep stages measured in the unary budget \((a+1)(u+1)\); the inequality itself is a statement of pure natural-number arithmetic.
Given a clock witness machine \(W\) over the Boolean alphabet and a parameter \(a \in \mathbb {N}\), the length-only schedule is a machine over the schedule alphabet with \(W.k + 3\) work tapes (the clock’s tapes plus a binary budget tape, a unary macrostep counter, and a sweep guide) whose states are the schedule phases over \(W\)’s state set. It first runs the input-masked clock, recording the clock’s emitted bits; copying and rescanning branch only on blank versus nonblank. Each successful decrement of the captured fixed-width binary budget appends \(a + 1\) unary markers, after an initial batch of \(a + 1\) markers, so the intended unary budget is \((a+1)(T(n)+1)\) where \(T(n)\) is the value emitted by the clock. The guide has radius three cells per unary marker; every macrostep starts at the guide’s marked origin, traverses the entire guide in both directions, and returns to the origin before advancing the unary counter, halting when the markers are exhausted. All of these actions are independent of the source decider and of all decorated data tapes.
Let \(W\) be a binary machine and \(a\) a padding constant, and consider the length-only schedule controller built from them, whose first administrative tape is the budget tape. If the budget tape currently reads a bit symbol \(b\), then the transition out of the borrow control phase with carry \(c\) is the single-lane action on the budget tape that writes the bit \(b \oplus c\), moves right, and enters the borrow phase with carry \(c \wedge \lnot b\).
Let \(W\) be a binary machine, \(a\) a padding constant, and let a bit word \(w\) be laid out as bit symbols on the budget lane of the schedule controller, starting in the borrow phase with carry \(c\). Running the controller for exactly \(|w|\) steps carries out the fixed-width borrow pass on \(w\): the phase becomes borrow with the resulting borrow flag, the rewritten word is moved in reverse into the processed zone, the head has advanced by \(|w|\), and every other tape content and head position, including copied input data in later decorations, is unchanged.
Let \(W\) be a machine over the Boolean alphabet, let \(a, b \in \mathbb {N}\), and let \(T \colon \mathbb {N}\to \mathbb {N}\) satisfy \(n \le T(n)\) for every \(n\). Suppose that on every Boolean input \(x\) the machine \(W\) halts within \(b\, (T(\left\lvert x\right\rvert )+1)\) steps with output exactly the list of (little-endian) binary digits of \(T(\left\lvert x\right\rvert )\). Then for every Boolean input \(x\), the length-only schedule machine built from \(W\) and the padding parameter \(a\), started on the embedding of \(x\) into the schedule alphabet (each bit sent to its tagged bit symbol), halts — through its clock capture, input reset, copy, fixed-width budget conversion, guide allocation, all macrosteps, and the final counter test — within a uniform quadratic budget: there exists a time
after which its control state is the halting state.
For every clock witness \(W\) over the Boolean alphabet and every parameter \(a \in \mathbb {N}\), the length-only schedule built from \(W\) and \(a\) is oblivious: on arbitrary equal-length inputs over the schedule alphabet, its input head and every clock, budget, counter, and guide head occupy the same position at every time. No computation or halting hypothesis is needed.
For every clock witness \(W\) over the Boolean alphabet, every parameter \(a \in \mathbb {N}\), every state \(q\) of the length-only schedule, every scanned input symbol, and every tuple of work-tape reads, the transition of the length-only schedule emits no output symbol. A decoration supplies the final answer at the schedule’s final halting transition.
Let \(W\) compute on every input \(x\) the binary expansion of \(T(|x|)\) within \(b\, (T(|x|)+1)\) steps, and suppose \(n \le T(n)\) for all \(n\). Then for every word \(x\) there is a time \(t\) at which the schedule controller built from \(W\) and \(a\), run on the embedded input, is in the counter-test phase with unary head at \(1\), guide head at the origin, unary counter tape with \((a+1)(T(|x|)+1)\) markers, and guide tape of radius \(3(a+1)(T(|x|)+1)\).
Let \(M\) be a binary machine with \(k\) work tapes, let \(q\) be an optional control state of \(M\), and let a vector of \(k+1\) payloads be given (one per work tape and one for the virtual input). If \(q\) is absent (the machine has halted), the selected action is the identity: no writes, no output, no head movement, and a halted successor state. If \(q\) is a live state, the selected action is the transition of \(M\) from \(q\) on the bit carried by the virtual-input payload and the bits carried by the work-tape payloads.
For every binary machine \(M\) and every configuration \(c\) of \(M\) on an input \(x\), the action selected by the data controller from the control state of \(c\) and the head-centered payloads at offset \(0\) equals the totalized action of \(M\) at \(c\).
Let \(M\) be a binary machine with \(k\) work tapes, let \(q\) be an optional control state, and let a vector of \(k+1\) payload reads be given. For each work tape the virtual head direction is the movement prescribed by the action selected from \(q\) and the reads; the virtual input head direction is that action’s input movement, clamped to stationary when the input payload carries the left-boundary tag and the movement is leftward, or the right-boundary tag and the movement is rightward. These directions select shifted payloads and never move a physical head.
For every binary machine \(M\) and every configuration \(c\) of \(M\) on an input \(x\), the vector of virtual head directions selected by the controller from the control state of \(c\) and the head-centered payloads at offset \(0\) equals the shift vector of the totalized action of \(M\) at \(c\); in particular native input clamping and source work head movement are reproduced exactly.
Let \(W\) and \(M\) be binary machines and \(a\) a padding constant. Given the current phase of the schedule controller, a data state for \(M\), the symbol under the input head, the reads of the controller’s tapes, and the reads of the \(k+1\) data tapes, this transducer returns an updated data state, an optional write for each data tape, and an optional output symbol. During the copy phases the input bit is written with explicit boundary payloads; at a marked origin with a positive counter test the source transition is selected and its writes performed; the forward sweep caches left neighbors; the backward sweep shifts each virtual tape according to its selected direction; the return to the origin installs the source successor state (possibly halted); and only the final failed unary-counter test emits the stored answer bit.
A payload is a pair consisting of an optional bit (the represented cell content, absent for a blank cell) and a tag from a three-element set: tag \(0\) marks an ordinary cell, tag \(1\) the left virtual input boundary, and tag \(2\) the right virtual input boundary. Payloads describe the cells of virtual tapes whose logical head sits at a fixed marked origin, so that moving a virtual head shifts the represented tape while the physical schedule stays fixed.
A finite type of control phases, parameterized by a set \(S\) (the states of the clock witness) and a natural number \(a\). It comprises an initial phase; a clock phase carrying a state of \(S\); phases for rescanning the input, copying and reversing the clock’s output, decrementing the binary budget (an append counter ranging over \(\{ 0, \dots , a+1\} \) and a borrow phase carrying a carry bit), and building the unary counter; layout phases for the sweep guide, each carrying an allocation phase in \(\{ 0, 1, 2\} \), together with edge and reset phases; and the sweep phases (start, macrostep check, seek left, forward, backward, return to centre). The source decider’s state is absent from this type.
A finite logical alphabet in which schedule symbols and data cells occupy disjoint parts. Its symbols are: a tagged bit (\(\mathsf{bit}\, b\) for \(b\) Boolean), an origin marker, a unary unit marker, an interior marker, a tagged edge marker (\(\mathsf{edge}\, b\)), and a data cell carrying two payloads, each a pair of a possibly-blank bit and an allocation phase in \(\{ 0, 1, 2\} \); in a data cell the second payload is the cached left neighbour.
The class \(\mathsf{P}\) of polynomial-time decidable languages over the binary alphabet: a language \(L \subseteq \{ 0,1\} ^{*}\) belongs to \(\mathsf{P}\) when there is a natural number \(c\) such that \(L \in \mathrm{DTIME}(n \mapsto n^c + 1)\), i.e. some deterministic finite Turing machine decides \(L\) within \(a \cdot (n^c + 1)\) steps on every input of length \(n\), for some constant \(a\). In symbols, \(\mathsf{P} = \bigcup _{c \in \mathbb {N}} \mathrm{DTIME}(n^c + 1)\).
Let \(L \subseteq \{ 0,1\} ^{*}\) be an \(\mathsf{NP}\)-hard language, i.e. every language in \(\mathsf{NP}\) is polynomial-time Karp reducible to \(L\). If \(L \in \mathsf{P}\), then \(\mathsf{P} = \mathsf{NP}\) [AB09, Theorem 2.8.2].
Every language \(L \subseteq \{ 0,1\} ^{*}\) decidable in polynomial time belongs to \(\mathsf{NP}\); that is, for every \(L \in \mathsf{P}\) there exist \(C, c \in \mathbb {N}\) and a language \(V \in \mathsf{P}\) such that for every binary string \(x\), \(x \in L\) if and only if some \(u \in \{ 0,1\} ^{*}\) with \(|u| = C \cdot (|x|+1)^c\) has \(xu \in V\).
Every language \(L \subseteq \{ 0,1\} ^{*}\) in \(\mathsf{P}\) (decidable in polynomial time) lies both in \(\mathsf{NP}\) and in coNP, where coNP is the class of languages whose complement is in \(\mathsf{NP}\); that is, \(\mathsf{P}\) is contained in the intersection of \(\mathsf{NP}\) and coNP.
For \(C, c \in \mathbb {N}\) and a language \(V \subseteq \{ 0,1\} ^{*}\), the language \(V^{\mathrm{pad}}_{C,c}\) consists of the binary strings \(y\) for which there is a least \(n \le |y|\) with \(n + (C+1)(n+1)^c = |y|\), and, writing \(y = xv\) with \(|x| = n\), the suffix \(v\) contains a \(1\), the prefix \(u\) of \(v\) preceding its last \(1\) satisfies \(|u| \le C(n+1)^c\), and \(\langle x, u\rangle \in V\). Here \(\langle x, u\rangle \) is the self-delimiting pair encoding: \(x\) with every bit doubled, then the separator \(01\), then \(u\) verbatim.
Let \(C, c \in \mathbb {N}\) and \(V \subseteq \{ 0,1\} ^{*}\), and let \(V^{\mathrm{pad}}_{C,c}\) be the set of binary strings \(y\) such that, for the least \(n\) with \(n + (C+1)(n+1)^c = |y|\) (which must exist), the suffix \(v\) of \(y\) after its first \(n\) bits contains a \(1\), and the prefix \(u\) of \(v\) preceding its last \(1\) satisfies \(|u| \le C(n+1)^c\) and \(\langle x, u\rangle \in V\), where \(x\) is the first \(n\) bits of \(y\) and \(\langle x, u\rangle \) is the self-delimiting pair encoding (\(x\) with every bit doubled, then \(01\), then \(u\)). Then for all binary strings \(x\) and \(v\) with \(|v| = (C+1)(|x|+1)^c\), the concatenation \(xv\) lies in \(V^{\mathrm{pad}}_{C,c}\) if and only if \(v\) contains a \(1\) and the prefix \(u\) of \(v\) preceding its last \(1\) satisfies \(|u| \le C(|x|+1)^c\) and \(\langle x, u\rangle \in V\).
Let \(C, c \in \mathbb {N}\) and let \(V \subseteq \{ 0,1\} ^{*}\) be a language in \(\mathsf{P}\). Consider the language \(V^{\mathrm{pad}}_{C,c}\) of binary strings \(y\) such that, for the least \(n\) with \(n + (C+1)(n+1)^c = |y|\) (which must exist), the suffix \(v\) of \(y\) after its first \(n\) bits contains a \(1\), and the prefix \(u\) of \(v\) preceding its last \(1\) satisfies \(|u| \le C(n+1)^c\) and \(\langle x, u\rangle \in V\), where \(x\) is the first \(n\) bits of \(y\) and \(\langle x, u\rangle \) is the self-delimiting pair encoding (\(x\) with every bit doubled, then \(01\), then \(u\)). Then \(V^{\mathrm{pad}}_{C,c} \in \mathsf{P}\).
Let \(C, c \in \mathbb {N}\) and \(V \subseteq \{ 0,1\} ^{*}\), and let \(V^{\mathrm{pad}}_{C,c}\) be the set of binary strings \(y\) such that, for the least \(n\) with \(n + (C+1)(n+1)^c = |y|\) (which must exist), the suffix \(v\) of \(y\) after its first \(n\) bits contains a \(1\), and the prefix \(u\) of \(v\) preceding its last \(1\) satisfies \(|u| \le C(n+1)^c\) and \(\langle x, u\rangle \in V\), where \(x\) is the first \(n\) bits of \(y\) and \(\langle x, u\rangle \) is the self-delimiting pair encoding (\(x\) with every bit doubled, then \(01\), then \(u\)). Then for every binary string \(x\) (even one containing \(1\)s), the string \(x\, 0^{(C+1)(|x|+1)^c}\) does not lie in \(V^{\mathrm{pad}}_{C,c}\).
Let \(C, c \in \mathbb {N}\) and \(V \subseteq \{ 0,1\} ^{*}\), and let \(V^{\mathrm{pad}}_{C,c}\) be the set of binary strings \(y\) such that, for the least \(n\) with \(n + (C+1)(n+1)^c = |y|\) (which must exist), the suffix \(v\) of \(y\) after its first \(n\) bits contains a \(1\), and the prefix \(u\) of \(v\) preceding its last \(1\) satisfies \(|u| \le C(n+1)^c\) and \(\langle x, u\rangle \in V\), where \(x\) is the first \(n\) bits of \(y\) and \(\langle x, u\rangle \) is the self-delimiting pair encoding (\(x\) with every bit doubled, then \(01\), then \(u\)). If \(y\) is a binary string such that no \(n \le |y|\) satisfies \(n + (C+1)(n+1)^c = |y|\), then \(y \notin V^{\mathrm{pad}}_{C,c}\).
Let \(C, c \in \mathbb {N}\) and \(V \subseteq \{ 0,1\} ^{*}\), and let \(V^{\mathrm{pad}}_{C,c}\) be the set of binary strings \(y\) such that, for the least \(n\) with \(n + (C+1)(n+1)^c = |y|\) (which must exist), the suffix \(v\) of \(y\) after its first \(n\) bits contains a \(1\), and the prefix \(u\) of \(v\) preceding its last \(1\) satisfies \(|u| \le C(n+1)^c\) and \(\langle x, u\rangle \in V\), where \(x\) is the first \(n\) bits of \(y\) and \(\langle x, u\rangle \) is the self-delimiting pair encoding (\(x\) with every bit doubled, then \(01\), then \(u\)). Let \(x, u\) be binary strings and \(k \in \mathbb {N}\) with \(|u\, 1\, 0^k| = (C+1)(|x|+1)^c\) and \(C(|x|+1)^c \lt |u|\). Then \(x\, u\, 1\, 0^k \notin V^{\mathrm{pad}}_{C,c}\).
Let \(C, c \in \mathbb {N}\) and \(V \subseteq \{ 0,1\} ^{*}\), and let \(V^{\mathrm{pad}}_{C,c}\) be the set of binary strings \(y\) such that, for the least \(n\) with \(n + (C+1)(n+1)^c = |y|\) (which must exist), the suffix \(v\) of \(y\) after its first \(n\) bits contains a \(1\), and the prefix \(u\) of \(v\) preceding its last \(1\) satisfies \(|u| \le C(n+1)^c\) and \(\langle x, u\rangle \in V\), where \(x\) is the first \(n\) bits of \(y\) and \(\langle x, u\rangle \) is the self-delimiting pair encoding (\(x\) with every bit doubled, then \(01\), then \(u\)). Then for every binary string \(x\), there exists \(v\) with \(|v| = (C+1)(|x|+1)^c\) and \(xv \in V^{\mathrm{pad}}_{C,c}\) if and only if there exists \(u\) with \(|u| \le C(|x|+1)^c\) and \(\langle x, u\rangle \in V\).
For \(C, c \in \mathbb {N}\) and a language \(V \subseteq \{ 0,1\} ^{*}\), the language \(V^{\mathrm{pair}}_{C,c}\) consists of the binary strings \(y\) that the aligned pair parser decodes as a pair \((x, u)\) — that is, \(y = \langle x, u\rangle \) is a well-formed self-delimiting pair encoding, consisting of \(x\) with every bit doubled, then the separator \(01\), then \(u\) verbatim — such that \(|u| = C(|x|+1)^c\) and the concatenation \(xu\) lies in \(V\).
Let \(C, c \in \mathbb {N}\) and \(V \subseteq \{ 0,1\} ^{*}\). The aligned pair parser reads a binary string in two-bit blocks from the left, taking \(00\) and \(11\) as the bits \(0\) and \(1\) of a first component until the first block \(01\), after which the remaining suffix is the second component; it fails on every other string. Let \(V^{\mathrm{pair}}_{C,c}\) be the set of strings that this parser decodes as a pair \((x, u)\) with \(|u| = C(|x|+1)^c\) and \(xu \in V\). If \(y\) is a binary string on which the parser fails, then \(y \notin V^{\mathrm{pair}}_{C,c}\).
Let \(C, c \in \mathbb {N}\) and let \(V \subseteq \{ 0,1\} ^{*}\) be a language in \(\mathsf{P}\). Then the language of binary strings \(y\) that are well-formed self-delimiting pair encodings \(y = \langle x, u\rangle \) (\(x\) with every bit doubled, then the separator \(01\), then \(u\) verbatim) with \(|u| = C(|x|+1)^c\) and \(xu \in V\) also belongs to \(\mathsf{P}\).
Let \(C, c \in \mathbb {N}\) and \(V \subseteq \{ 0,1\} ^{*}\), write \(\langle x, u\rangle \) for the self-delimiting pair encoding of binary strings \(x\) and \(u\) (\(x\) with every bit doubled, then the separator \(01\), then \(u\)), and let \(V^{\mathrm{pair}}_{C,c}\) be the set of strings that the aligned pair parser decodes as a pair \((x, u)\) with \(|u| = C(|x|+1)^c\) and \(xu \in V\). Then for all binary strings \(x\) and \(u\), \(\langle x, u\rangle \in V^{\mathrm{pair}}_{C,c}\) if and only if \(|u| = C(|x|+1)^c\) and \(xu \in V\).
Let \(\langle a,b\rangle \) denote the self-delimiting pair encoding of bit strings \(a\) and \(b\) (every bit of \(a\) doubled, then \(01\), then \(b\)). For all bit strings \(a, b, c\), appending \(c\) to \(\langle a,b\rangle \) appends it to the second component: \(\langle a,b\rangle \, c = \langle a, b\, c\rangle \), where juxtaposition denotes concatenation.
Let \(\langle a,b\rangle \) denote the self-delimiting pair encoding of bit strings \(a\) and \(b\) (every bit of \(a\) doubled, then \(01\), then \(b\)). The total first projection \(\pi _1\) sends a bit string \(z\) to \(a\) when \(z = \langle a,b\rangle \) is a pair encoding, and to the empty word when \(z\) is not of this form.
Let \(\langle a,b\rangle \) denote the self-delimiting pair encoding of bit strings \(a\) and \(b\) (every bit of \(a\) doubled, then \(01\), then \(b\)), and let \(\pi _1\) be the total first projection, which returns the first component of a pair encoding and the empty word on any other bit string. Then \(\pi _1(\langle a,b\rangle ) = a\) for all bit strings \(a, b\).
Let \(g\) be a function from bit strings to bit strings, and let \(\langle a,b\rangle \) denote the self-delimiting pair encoding of bit strings \(a\) and \(b\) (every bit of \(a\) doubled, then \(01\), then \(b\)). The threaded payload map of \(g\) sends a bit string \(z\) to \(\langle a, g(b)\rangle \) when \(z = \langle a,b\rangle \) is a pair encoding, carrying the first component unchanged, and to the empty word when \(z\) is not of this form.
Let \(g\) be a function from bit strings to bit strings, let \(\langle a,b\rangle \) denote the self-delimiting pair encoding of bit strings \(a,b\) (every bit of \(a\) doubled, then \(01\), then \(b\)), and let \(\sigma _g\) be the threaded payload map of \(g\), which decodes a bit string \(z\) as a pair \(\langle a,b\rangle \) and outputs \(\langle a, g(b)\rangle \), or outputs the empty word if \(z\) is not a pair encoding. Then \(\sigma _g(\langle a,b\rangle ) = \langle a, g(b)\rangle \) for all bit strings \(a\) and \(b\).
Let \(\langle a,b\rangle \) denote the self-delimiting pair encoding of bit strings \(a\) and \(b\) (every bit of \(a\) doubled, then \(01\), then \(b\)). The total second projection \(\pi _2\) sends a bit string \(z\) to \(b\) when \(z = \langle a,b\rangle \) is a pair encoding, and to the empty word when \(z\) is not of this form.
Let \(\langle a,b\rangle \) denote the self-delimiting pair encoding of bit strings \(a\) and \(b\) (every bit of \(a\) doubled, then \(01\), then \(b\)), and let \(\pi _2\) be the total second projection, which returns the second component of a pair encoding and the empty word on any other bit string. Then \(\pi _2(\langle a,b\rangle ) = b\) for all bit strings \(a, b\).
Let \(x\) be a binary word. For every \(t \le |x|\), running the three-state palindrome machine for \(t\) steps from its initial configuration on input \(x\) reaches the copy-state configuration with input head at cell \(t+1\), exactly the first \(t\) bits of \(x\) copied onto the work tape (every other work cell still blank), and work head at cell \(t\).
Let \(x\) be a binary word of length \(n\). From the configuration of the three-state palindrome machine on input \(x\) in copy state with the input head on the right boundary blank (cell \(n+1\)), all \(n\) bits copied, and work head at cell \(n\), one transition switches to the rewind state and moves the input head left to cell \(n\), leaving the work tape and the work head unchanged.
Let \(x\) be a binary word and \(t \lt |x|\). From the configuration of the three-state palindrome machine on input \(x\) in copy state with input head at cell \(t+1\) (scanning \(x_t\)), the first \(t\) bits copied to the work tape, and work head at cell \(t\), one transition writes \(x_t\) at the work head and advances both heads: it yields the copy-state configuration with input head at \(t+2\), the first \(t+1\) bits copied, and work head at \(t+1\).
The language of binary palindromes — the set of words \(x \in \{ 0,1\} ^*\) with \(x^{R} = x\) — lies in \(\mathrm{DTIME}(n+1)\): there exist a constant \(c\) and a multi-tape Turing machine over the binary alphabet that, on every input of length \(n\), halts within \(c\, (n+1)\) steps and outputs true exactly when the input is a palindrome, and false otherwise.
The language of binary palindromes — the set of words \(x \in \{ 0,1\} ^*\) equal to their own reversal — belongs to the class \(\mathsf{P}\) of languages decidable by a multi-tape Turing machine in time polynomial in the input length.
Let \(x\) be a binary word of length \(n\). For every \(i \le n\), running the three-state palindrome machine for \(i\) steps from the rewind-state configuration on input \(x\) with input head at cell \(i\), all \(n\) bits copied, and work head at cell \(n\) brings the input head to the left boundary cell \(0\), with the state, work tape, and work head unchanged.
Let \(x\) be a binary word of length \(n\) and \(i \lt n\). From the configuration of the three-state palindrome machine on input \(x\) in rewind state with input head at cell \(i+1\), all \(n\) bits copied, and work head at cell \(n\), one transition moves the input head left to cell \(i\) and changes nothing else: the state, the copied work tape, and the work head position are preserved.
Let \(x\) be a binary word of length \(n\), and let \(r\) and \(i \le n\) satisfy \(n = i + r\). Running the three-state palindrome machine for \(r+1\) steps from the test-state configuration on input \(x\) with input head at cell \(i+1\), all \(n\) bits copied, and work head at the integer cell \(n-1-i\) produces a halted configuration whose output is the single bit true if every opposing pair of \(x\) at or after position \(i\) agrees (the \(j\)-th bit of \(x^{R}\) equals the \(j\)-th bit of \(x\) for all \(i \le j \lt n\)), and false otherwise.
Let \(x\) be a binary word of length \(n\) and \(i \lt n\). In the test-state configuration of the palindrome machine on input \(x\) with input head at cell \(i+1\), all \(n\) bits copied, and work head at the integer cell \(n-1-i\), the symbol under the work head is the \(i\)-th bit of the reversal \(x^{R}\).
Let \(x\) be a binary word of length \(n\). From the configuration of the three-state palindrome machine on input \(x\) in rewind state with the input head on the left boundary blank (cell \(0\)), all \(n\) bits copied, and work head at cell \(n\), one transition switches to the test state, moves the input head right to cell \(1\), and moves the work head left to the integer cell \(n-1\) (which is \(-1\) when \(x\) is empty), leaving the work tape unchanged.
Let \(x\) be a binary word of length \(n\) and \(i \lt n\), and suppose the \(i\)-th bit of the reversal \(x^{R}\) equals the \(i\)-th bit of \(x\). From the test-state configuration of the palindrome machine on input \(x\) with input head at cell \(i+1\), all \(n\) bits copied, and work head at the integer cell \(n-1-i\), one transition advances the opposing heads: the input head moves right to cell \(i+2\) and the work head moves left to cell \(n-1-(i+1)\), with state and work tape unchanged.
For a binary word \(x\) of length \(n\), a control state \(q \in \{ 0,1,2\} \), an input head position \(p \in \{ 0,\dots ,n+1\} \) (cells \(1\) through \(n\) hold the input bits, cells \(0\) and \(n+1\) are blank boundary cells), an integer work head position \(w\), and a time bound \(t\), this is the machine configuration in live state \(q\) with input head at \(p\), work tape holding exactly the first \(t\) bits of \(x\) in cells \(0,\dots ,t-1\) (all other cells blank), work head at \(w\), and empty output. These are the configurations appearing in the three phase invariants.
For a binary word \(x\) and an index \(i\), the predicate asserting that every opposing pair at or after position \(i\) agrees: for all \(j\) with \(i \le j \lt |x|\), the \(j\)-th bit of the reversal \(x^{R}\) equals the \(j\)-th bit of \(x\).
Let \(x\) be a binary word and \(i \lt |x|\) an index at which the \(i\)-th bit of the reversal \(x^{R}\) equals the \(i\)-th bit of \(x\). Then all opposing pairs of \(x\) at or after position \(i\) agree if and only if all opposing pairs at or after position \(i+1\) agree.
Given a binary word \(x\), a time bound \(t \in \mathbb {N}\), and an integer cell index \(z\), this is the partial tape content whose cell \(z\) holds the bit \(x_z\) when \(0 \le z \lt t\) (and \(z\) is within \(x\)), and is blank otherwise. It describes a work tape onto which exactly the first \(t\) bits of \(x\) have been copied.
Let \(x\) be a binary word and \(t \lt |x|\). Updating the partial tape that holds the first \(t\) bits of \(x\) (in cells \(0,\dots ,t-1\), blank elsewhere) by writing the bit \(x_t\) into cell \(t\) yields exactly the partial tape holding the first \(t+1\) bits of \(x\).
A concrete Turing machine over the binary alphabet with one work tape and three control states \(\{ 0,1,2\} \), given by an explicit transition table. State \(0\) (copy) copies each input bit to the work tape while moving both heads right; upon reading the blank past the input it switches to state \(1\) (rewind), which moves the input head back left to the left boundary; state \(2\) (test) then moves the input head right and the work head left in unison, comparing the two scanned symbols, emitting false and halting on any mismatch and emitting true when the input is exhausted with all pairs matched.
Let \(A\) be a finite alphabet with decidable equality, \(e\) an embedding of the Booleans into \(A\), \(x\) a binary word, and \(c\) a configuration of a \(k\)-tape machine over \(A\) on the input obtained by embedding \(x\) symbolwise. The encoded configuration over the Boolean alphabet has the same control state, the same numerical input-head position, work tapes indexed by pairs \((i, a)\) whose contents are track \(a\) of the transverse encoding of tape \(i\) of \(c\), head positions on every track of block \(i\) equal to \(c\)’s head position on tape \(i\), and output stream obtained from \(c\)’s by decoding each symbol to a bit.
Let \(P\) be a machine over a finite alphabet \(A\), \(e\) an embedding of the Booleans into \(A\), and \(x\) a binary word. Encoding the initial configuration of \(P\) on the symbolwise embedding of \(x\) yields exactly the initial configuration of the transverse binary realization of \(P\) on \(x\).
Let \(c\) be a configuration of a machine over a finite alphabet \(A\) on the symbolwise embedding of a binary word \(x\) under an embedding \(e\) of the Booleans into \(A\). Mapping the (possibly blank) input bit scanned by the encoded configuration through \(e\) recovers exactly the (possibly blank) input symbol scanned by \(c\).
Let \(c\) be a configuration of a \(k\)-tape machine over a finite alphabet \(A\) on the symbolwise embedding of a binary word. For every logical tape index \(i \lt k\), decoding the family of bits scanned on the tracks of block \(i\) in the encoded configuration yields exactly the symbol scanned by \(c\) on tape \(i\).
An injection from possibly-blank symbols over an alphabet \(A\) (with decidable equality) into families of possibly-blank Booleans indexed by \(A\): a nonblank symbol \(a\) is sent to the indicator family whose track \(j \in A\) holds the bit \([\, j = a\, ]\), and the blank symbol is sent to the all-blank family. A symbol is thus stored transversely, one track per element of \(A\), so that all tracks of a logical cell share one head coordinate and a logical read or write takes exactly one physical transition.
A total (noncomputable) decoding function from families of possibly-blank Booleans indexed by \(A\) back to possibly-blank symbols of \(A\): on the image of the transverse encoding it returns the unique preimage, and on malformed families it returns a fixed chosen value. Initialized simulations only ever read encoded blocks.
Let \(P\) be a machine with \(k\) work tapes over a finite alphabet \(A\) and let \(e\) be an embedding of the Booleans into \(A\). The transverse binary realization is a machine over the Boolean alphabet with the same state set and one work tape per pair \((i, a) \in \{ 0, \dots , k-1\} \times A\) (so \(k \cdot \left\lvert A\right\rvert \) tapes in all). At each step it decodes the scanned tracks of every logical tape to a symbol of \(A\), maps the scanned input bit through \(e\), consults \(P\)’s transition, re-encodes the chosen write transversely on the corresponding block of tracks, moves all tracks of each block by \(P\)’s move for that tape, and decodes any emitted symbol to a bit. Every logical step takes exactly one physical step.
Let \(P\) be a machine over a finite alphabet \(A\), \(e\) an embedding of the Booleans into \(A\), \(x\) and \(w\) binary words, and \(t \in \mathbb {N}\). If, started on the symbolwise embedding of \(x\), the machine \(P\) has halted after \(t\) steps with output exactly the symbolwise embedding of \(w\), then the transverse binary realization of \(P\), started on \(x\), has halted after the same \(t\) steps with output exactly \(w\).
Let \(P\) be a machine over a finite alphabet \(A\) and \(e\) an embedding of the Booleans into \(A\). If \(P\) is oblivious, then the transverse binary realization of \(P\) is oblivious — at all intermediate physical times and at all times after halting.
Let \(P\) be a machine over a finite alphabet \(A\), \(e\) an embedding of the Booleans into \(A\), \(x\) a binary word, and \(t \in \mathbb {N}\). The configuration of the transverse binary realization of \(P\) after \(t\) steps from its initial configuration on \(x\) equals the encoding of the configuration of \(P\) after \(t\) steps from its initial configuration on the symbolwise embedding of \(x\) — with no intermediate coding states and no variable-duration memory accesses.
Let \(P\) be a machine over a finite alphabet \(A\), \(e\) an embedding of the Booleans into \(A\), and \(c\) a configuration of \(P\) on the symbolwise embedding of a binary word \(x\). Then one step of the transverse binary realization of \(P\) at the encoded configuration of \(c\) equals the encoded configuration of one step of \(P\) at \(c\): both the read and the write of a logical cell happen simultaneously across its fixed collection of tracks, and the halted case is absorbing on both sides.
Let \(c\) be a configuration of a binary machine, \(\alpha \) an action, \(z\) a coordinate, and \(n\) a width. Folding the backward transducer with the shift directions of \(\alpha \), seeded with the written payloads of \(c\) and \(\alpha \) at \(z + n\), over the reversed cached zone of written payloads starting at \(z\) yields the written payloads at \(z\) together with the reversed uncached zone of the head-centered payloads of the configuration obtained by applying \(\alpha \) to \(c\); the two finite transductions thus realize one source action on the whole fixed layout.
Given per-tape shift directions, a right-neighbor register, and a row, the backward transducer outputs the row’s payloads as the new register and rewrites each tape’s entry to its cached left neighbor, its current payload, or the incoming right register according as that tape’s direction is leftward, stationary, or rightward, clearing the cache to blank; all physical heads move left while it runs.
For every payload family \(f\), every vector of directions \(d\), and every coordinate \(z\), applying the backward transducer with directions \(d\) and register the payloads at \(z + 1\) to the cached row at \(z\) produces the payloads at \(z\) as the new register together with the uncached row at \(z\) of the shifted family sending coordinate \(z\) and tape \(i\) to the value of \(f\) at \(z\) plus the direction of \(i\).
For every payload family \(f\), every vector of directions \(d\), every coordinate \(z\), and every length \(n\), folding the backward transducer with directions \(d\) and initial register the payloads at \(z + n\) over the reversed cached zone of \(n\) rows starting at \(z\) ends with register the payloads at \(z\) and produces the reversed uncached zone of the family shifted per tape by \(d\); every row is shifted as prescribed.
For every payload family, coordinate \(z\), and length \(n\), folding the forward transducer with initial register the payloads at \(z - 1\) over the uncached zone of \(n\) rows starting at \(z\) ends with register the payloads at \(z + n - 1\) and produces the cached zone of the same \(n\) rows; all left-neighbor caches are installed.
Given a family of payloads indexed by an integer coordinate and a tape index, a coordinate \(z\), and a caching flag, the row at \(z\) assigns to each tape the pair of its payload at \(z\) together with, when the flag is set, its payload at \(z - 1\) as cached left neighbor, and the blank payload otherwise.
Let \(c, C, q \in \mathbb {N}\), \(x\) a bit string, \(h \colon \{ 0,\dots ,c\} \to \mathbb {Z}\), \(o\) an output word and \(i \in \{ 0,\dots ,c\} \). One step of \(\mathcal{P}_{c,C}\) from \(\Gamma _q(\mathrm{advance}_i, h, o)\) yields \(\Gamma _q(\mathrm{loop}_i, h[i \mapsto h(i)+1], o)\); here \(\mathcal{P}_{c,C}\) is the unary box-enumeration machine with \(c+1\) work tapes, which writes \(1^{|x|+1}\) on every work tape and then runs \(c+1\) nested unary loops, emitting \(C\) ones per innermost iteration, and \(\mathrm{loop}_i\), \(\mathrm{rewind}_i\), \(\mathrm{advance}_i\), \(\mathrm{emit}_j\) are its control states, \(\Gamma _q(s,h,o)\) denotes its configuration on input \(x\) in control state \(s\), with the input head on the blank cell right of the input, every work tape holding \(1^q\) in cells \(0,\dots ,q-1\) and blanks elsewhere, work head \(k\) at position \(h(k)\), and output \(o\), and \(h[i \mapsto v]\) is \(h\) with its \(i\)-th entry replaced by \(v\).
Let \(c, C \in \mathbb {N}\) and let \(x\) be a bit string. For every \(i \le |x|\), running \(\mathcal{P}_{c,C}\) for \(i\) steps from its initial configuration on \(x\) yields the configuration in the copy state with the input head on input position \(i\) (counting from \(0\)), every work tape holding \(1^i\) in cells \(0,\dots ,i-1\) and blanks elsewhere, every work head at cell \(i\), and empty output; here \(\mathcal{P}_{c,C}\) is the unary box-enumeration machine with \(c+1\) work tapes, which writes \(1^{|x|+1}\) on every work tape and then runs \(c+1\) nested unary loops, emitting \(C\) ones per innermost iteration, and \(\mathrm{loop}_i\), \(\mathrm{rewind}_i\), \(\mathrm{advance}_i\), \(\mathrm{emit}_j\) are its control states.
Let \(c, C, q \in \mathbb {N}\), \(x\) a bit string and \(h \colon \{ 0,\dots ,c\} \to \mathbb {Z}\). For every \(j \le C\) and every output word \(o\), running \(\mathcal{P}_{c,C}\) for exactly \(j+1\) steps from \(\Gamma _q(\mathrm{emit}_j, h, o)\) yields \(\Gamma _q(\mathrm{advance}_0, h, o\, 1^j)\); here \(\mathcal{P}_{c,C}\) is the unary box-enumeration machine with \(c+1\) work tapes, which writes \(1^{|x|+1}\) on every work tape and then runs \(c+1\) nested unary loops, emitting \(C\) ones per innermost iteration, and \(\mathrm{loop}_i\), \(\mathrm{rewind}_i\), \(\mathrm{advance}_i\), \(\mathrm{emit}_j\) are its control states, and \(\Gamma _q(s,h,o)\) denotes its configuration on input \(x\) in control state \(s\), with the input head on the blank cell right of the input, every work tape holding \(1^q\) in cells \(0,\dots ,q-1\) and blanks elsewhere, work head \(k\) at position \(h(k)\), and output \(o\).
Let \(c, C, q \in \mathbb {N}\) with \(q \gt 0\), \(x\) a bit string, \(i \le c\), \(h \colon \{ 0,\dots ,c\} \to \mathbb {Z}\) with \(h(k) = 0\) for every \(k \le i\), \(o\) an output word, and \(r, j \in \mathbb {N}\) with \(j + r = q\). Then running \(\mathcal{P}_{c,C}\) for exactly \(r\, (\kappa _{q,C}(i) + 2) + q + 2\) steps from \(\Gamma _q(\mathrm{loop}_i, h[i \mapsto j], o)\) yields \(\Gamma _q(\mathrm{advance}_{i+1}, h, o\, 1^{rCq^i})\), where \(\kappa _{q,C}(0) = C+1\) and \(\kappa _{q,C}(r'+1) = q(\kappa _{q,C}(r')+2)+q+2\); here \(\mathcal{P}_{c,C}\) is the unary box-enumeration machine with \(c+1\) work tapes, which writes \(1^{|x|+1}\) on every work tape and then runs \(c+1\) nested unary loops, emitting \(C\) ones per innermost iteration, and \(\mathrm{loop}_i\), \(\mathrm{rewind}_i\), \(\mathrm{advance}_i\), \(\mathrm{emit}_j\) are its control states, \(\Gamma _q(s,h,o)\) denotes its configuration on input \(x\) in control state \(s\), with the input head on the blank cell right of the input, every work tape holding \(1^q\) in cells \(0,\dots ,q-1\) and blanks elsewhere, work head \(k\) at position \(h(k)\), and output \(o\), and \(h[i \mapsto v]\) is \(h\) with its \(i\)-th entry replaced by \(v\).
Let \(c, C, q \in \mathbb {N}\), \(x\) a bit string, \(h \colon \{ 0,\dots ,c\} \to \mathbb {Z}\), \(o\) an output word and \(i \in \{ 0,\dots ,c\} \). For every \(j \le q\), running \(\mathcal{P}_{c,C}\) for exactly \(j+1\) steps from \(\Gamma _q(\mathrm{rewind}_i, h[i \mapsto j-1], o)\) yields \(\Gamma _q(\mathrm{advance}_{i+1}, h[i \mapsto 0], o)\); here \(\mathcal{P}_{c,C}\) is the unary box-enumeration machine with \(c+1\) work tapes, which writes \(1^{|x|+1}\) on every work tape and then runs \(c+1\) nested unary loops, emitting \(C\) ones per innermost iteration, and \(\mathrm{loop}_i\), \(\mathrm{rewind}_i\), \(\mathrm{advance}_i\), \(\mathrm{emit}_j\) are its control states, \(\Gamma _q(s,h,o)\) denotes its configuration on input \(x\) in control state \(s\), with the input head on the blank cell right of the input, every work tape holding \(1^q\) in cells \(0,\dots ,q-1\) and blanks elsewhere, work head \(k\) at position \(h(k)\), and output \(o\), and \(h[i \mapsto v]\) is \(h\) with its \(i\)-th entry replaced by \(v\).
Let \(c, C, q \in \mathbb {N}\) and let \(x\) be a bit string. For every \(j \le q\), running \(\mathcal{P}_{c,C}\) for exactly \(j+1\) steps from \(\Gamma _q(\mathrm{setup}, h, \varepsilon )\), where every work head is at position \(h(k) = j-1\), yields \(\Gamma _q(\mathrm{loop}_c, h_0, \varepsilon )\), where every work head is at \(h_0(k) = 0\); here \(\mathcal{P}_{c,C}\) is the unary box-enumeration machine with \(c+1\) work tapes, which writes \(1^{|x|+1}\) on every work tape and then runs \(c+1\) nested unary loops, emitting \(C\) ones per innermost iteration, and \(\mathrm{loop}_i\), \(\mathrm{rewind}_i\), \(\mathrm{advance}_i\), \(\mathrm{emit}_j\) are its control states, setup is its rewinding state, and \(\Gamma _q(s,h,o)\) denotes its configuration on input \(x\) in control state \(s\), with the input head on the blank cell right of the input, every work tape holding \(1^q\) in cells \(0,\dots ,q-1\) and blanks elsewhere, work head \(k\) at position \(h(k)\), and output \(o\).
Let \(c, C \in \mathbb {N}\) and let \(x\) be a bit string. After exactly \(2(|x|+1)\) steps from its initial configuration on \(x\), \(\mathcal{P}_{c,C}\) is in control state \(\mathrm{loop}_c\), with its input head on the blank cell right of the input, every work tape holding \(1^{|x|+1}\) in cells \(0,\dots ,|x|\) and blanks elsewhere, every work head at cell \(0\), and empty output; here \(\mathcal{P}_{c,C}\) is the unary box-enumeration machine with \(c+1\) work tapes, which writes \(1^{|x|+1}\) on every work tape and then runs \(c+1\) nested unary loops, emitting \(C\) ones per innermost iteration, and \(\mathrm{loop}_i\), \(\mathrm{rewind}_i\), \(\mathrm{advance}_i\), \(\mathrm{emit}_j\) are its control states.
For all \(c, C \in \mathbb {N}\), the unary box-enumeration machine \(\mathcal{P}_{c,C}\) (a finite binary-alphabet machine with \(c+1\) work tapes that runs \(c+1\) nested unary loops of side \(|x|+1\), emitting \(C\) ones per innermost iteration) computes the function \(x \mapsto 1^{C(|x|+1)^{c+1}}\) in time \(n \mapsto (C + 5(c+1) + 4)(n+1)^{c+1}\): on every input \(x\) it halts within \((C + 5(c+1) + 4)(|x|+1)^{c+1}\) steps with output exactly \(C(|x|+1)^{c+1}\) ones.
A function \(p : \mathbb {N} \to \mathbb {N}\) is polynomially bounded if there are constants \(C, c \in \mathbb {N}\) such that \(p(n) \le C(n+1)^c\) for every \(n \in \mathbb {N}\). The function \(p\) itself need be neither monotone nor computable; only the majorant \(C(n+1)^c\) is.
Let \(c, C, q \in \mathbb {N}\) and let \(x\) be a bit string. For a control state \(s\) of the unary box-enumeration machine \(\mathcal{P}_{c,C}\), head positions \(h \colon \{ 0,\dots ,c\} \to \mathbb {Z}\) and an output word \(o\), the loop configuration \(\Gamma _q(s,h,o)\) is the configuration of a \((c+1)\)-work-tape machine on input \(x\) that is in state \(s\), has its input head on the blank cell immediately right of the input, has every work tape holding \(1^q\) in cells \(0,\dots ,q-1\) and blanks elsewhere, has work head \(k\) at position \(h(k)\), and has output \(o\).
For natural numbers \(c\) and \(C\), this is the finite set of control states of the unary box-enumeration machine: a copy state, a setup state, a loop state \(\mathrm{loop}_i\) and a rewind state \(\mathrm{rewind}_i\) for each level \(i \in \{ 0,\dots ,c\} \), an advance state \(\mathrm{advance}_i\) for each \(i \in \{ 0,\dots ,c+1\} \), and an emission state \(\mathrm{emit}_j\) for each \(j \in \{ 0,\dots ,C\} \), recording how many output bits remain to be emitted.
Let \(c, C \in \mathbb {N}\), \(x\) a bit string and \(i \le |x|\). The copying configuration at stage \(i\) is the configuration of the unary box-enumeration machine \(\mathcal{P}_{c,C}\) (with \(c+1\) work tapes) on input \(x\) in the copy state, with the input head on input position \(i\) (counting from \(0\); the blank cell right of the input when \(i = |x|\)), every work tape holding \(1^i\) in cells \(0,\dots ,i-1\) and blanks elsewhere, every work head at cell \(i\), and empty output.
For \(q, C \in \mathbb {N}\), the function \(\kappa _{q,C} \colon \mathbb {N} \to \mathbb {N}\) is defined by \(\kappa _{q,C}(0) = C + 1\) and \(\kappa _{q,C}(r+1) = q\, (\kappa _{q,C}(r) + 2) + q + 2\); it is the exact number of steps of a full nest of \(r\) unary loops of side \(q\) that emit \(C\) bits per innermost iteration.
Let \(q, C \in \mathbb {N}\) with \(q \gt 0\), and let \(\kappa _{q,C} \colon \mathbb {N} \to \mathbb {N}\) be given by \(\kappa _{q,C}(0) = C + 1\) and \(\kappa _{q,C}(r+1) = q(\kappa _{q,C}(r) + 2) + q + 2\). Then \(\kappa _{q,C}(r) \le (C + 1 + 5r)\, q^r\) for every \(r \in \mathbb {N}\).
For \(c, C \in \mathbb {N}\), a work-tape index \(i \in \{ 0,\dots ,c\} \), a direction \(d \in \{ -1,0,1\} \) and a control state \(s\) of the unary box enumerator, this is the one-step action of a \((c+1)\)-work-tape binary machine that keeps the input head in place, writes nothing on any work tape, moves work head \(i\) by \(d\) while leaving every other work head in place, emits no output, and passes to state \(s\).
Let \(c, C, q \in \mathbb {N}\), \(x\) a bit string, \(s, s'\) control states of the unary box enumerator \(\mathcal{P}_{c,C}\), \(h \colon \{ 0,\dots ,c\} \to \mathbb {Z}\), \(o\) an output word, \(i \in \{ 0,\dots ,c\} \) and \(d \in \{ -1,0,1\} \). Applying the action that moves only work head \(i\) by \(d\) (writing nothing, emitting nothing, input head fixed) and enters \(s'\) to \(\Gamma _q(s,h,o)\) gives \(\Gamma _q(s', h[i \mapsto h(i)+d], o)\); here \(\Gamma _q(s,h,o)\) denotes its configuration on input \(x\) in control state \(s\), with the input head on the blank cell right of the input, every work tape holding \(1^q\) in cells \(0,\dots ,q-1\) and blanks elsewhere, work head \(k\) at position \(h(k)\), and output \(o\), and \(h[i \mapsto v]\) is \(h\) with its \(i\)-th entry replaced by \(v\).
For every \(q \in \mathbb {N}\), writing the symbol \(1\) into cell \(q\) of the tape contents \(\tau _q\) yields \(\tau _{q+1}\), where \(\tau _q \colon \mathbb {Z} \to \{ 0,1,\sqcup \} \) holds \(1\) on cells \(0,\dots ,q-1\) and the blank \(\sqcup \) elsewhere.
A function \(f : \{ 0,1\} ^* \to \{ 0,1\} ^*\) on bit strings is polynomial-time computable if there exist a Turing machine \(M\) over the binary alphabet with finitely many states and constants \(C, c \in \mathbb {N}\) such that, on every input \(x\) of length \(n\), \(M\) halts within \(C(n+1)^c\) steps with \(f(x)\) written on its output tape. These functions form the class \(\mathrm{FP}\) implicit throughout [AB09, ch. 2].
Let \(f\) and \(g\) be polynomial-time computable functions on bit strings (each computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then the map \(x \mapsto f(x)\, g(x)\), sending \(x\) to the concatenation of \(f(x)\) and \(g(x)\), is polynomial-time computable.
Let \(f, g : \{ 0,1\} ^* \to \{ 0,1\} ^*\) be polynomial-time computable, where a function is polynomial-time computable if some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\). Then the composition \(g \circ f\), mapping \(x\) to \(g(f(x))\), is also polynomial-time computable [AB09, proof of Theorem 2.8].
Let \(f : \{ 0,1\} ^* \to \{ 0,1\} ^*\) be polynomial-time computable, i.e. some finite binary Turing machine computes \(f\) within \(C_0(n+1)^{c_0}\) steps on every input of length \(n\), for some constants \(C_0, c_0 \in \mathbb {N}\). Then there are constants \(C, c \in \mathbb {N}\) such that \(|f(x)| \le C(|x|+1)^c\) for every bit string \(x\).
Let \(f\) and \(g\) be polynomial-time computable functions on bit strings (each computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then the map \(x \mapsto \langle f(x), g(x)\rangle \) is polynomial-time computable, where \(\langle a,b\rangle \) denotes the self-delimiting pair encoding of bit strings \(a\) and \(b\) (every bit of \(a\) doubled, then \(01\), then \(b\)).
Let \(g\) be a polynomial-time computable function on bit strings (computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then its threaded payload map, which sends a pair encoding \(\langle a,b\rangle \) (every bit of \(a\) doubled, then \(01\), then \(b\)) to \(\langle a, g(b)\rangle \) and every bit string that is not a pair encoding to the empty word, is also polynomial-time computable.
Let \(p\) and \(q\) be Boolean-valued predicates on bit strings such that the maps sending \(x\) to the one-bit words \(p(x)\) and \(q(x)\) are both polynomial-time computable (computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then the map sending \(x\) to the one-bit word \(p(x) \land q(x)\) is polynomial-time computable.
For every fixed bit string \(w\), the constant function \(x \mapsto w\) on bit strings is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
The identity function \(x \mapsto x\) on bit strings is polynomial-time computable: some finite binary Turing machine outputs \(x\) within \(C(n+1)^c\) steps on every input \(x\) of length \(n\), for some constants \(C, c \in \mathbb {N}\).
Let \(p\) be a Boolean-valued predicate on bit strings and let \(f, g\) be functions from bit strings to bit strings. Suppose that the map sending \(x\) to the one-bit word \(p(x)\) is polynomial-time computable, and that \(f\) and \(g\) are polynomial-time computable (each computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then the map sending \(x\) to \(f(x)\) if \(p(x)\) is true and to \(g(x)\) otherwise is polynomial-time computable.
Let \(\pi _1\) be the total first projection, which sends a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(a\) and any other bit string to the empty word. For every \(n \in \mathbb {N}\), the \(n\)-fold iterate \(\pi _1^{n} = \pi _1 \circ \cdots \circ \pi _1\) (the identity for \(n = 0\)), which extracts the root of a left-nested tuple, is polynomial-time computable.
Let \(\pi _1\) and \(\pi _2\) be the total projections, which send a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(a\) and to \(b\) respectively, and any other bit string to the empty word. Then the map sending a bit string \(z\) to the one-bit word that is \(1\) if \(|\pi _1(z)| = |\pi _2(z)|\) and \(0\) otherwise is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
Let \(\pi _1\) and \(\pi _2\) be the total projections, which send a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(a\) and to \(b\) respectively, and any other bit string to the empty word. Then the map sending a bit string \(z\) to the one-bit word that is \(1\) if \(|\pi _2(z)| \le |\pi _1(z)|\) and \(0\) otherwise is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
Let \(f\) be a function from bit strings to bit strings. If there are a finite binary Turing machine \(M\) and a constant \(a \in \mathbb {N}\) such that \(M\) computes \(f\) within \(a(n+1)\) steps on every input of length \(n\), then \(f\) is polynomial-time computable, i.e. some finite binary Turing machine computes \(f\) within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
Let \(\pi _1\) and \(\pi _2\) be the total projections, which send a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(a\) and to \(b\) respectively, and any other bit string to the empty word. Then the map \(z \mapsto \pi _1(z)\, \pi _2(z)\), sending \(z\) to the concatenation of its two components, is polynomial-time computable.
The total first projection \(\pi _1\), which sends a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(a\) and any other bit string to the empty word, is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
The total second projection \(\pi _2\), which sends a bit string of the form \(\langle a,b\rangle \) (the self-delimiting pair encoding: every bit of \(a\) doubled, then \(01\), then \(b\)) to \(b\) and any other bit string to the empty word, is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
Let \(\langle a,b\rangle \) denote the self-delimiting pair encoding of bit strings \(a\) and \(b\) (every bit of \(a\) doubled, then \(01\), then \(b\)), and let \(\pi _1\) and \(\pi _2\) be the total projections, which send \(\langle a,b\rangle \) to \(a\) and to \(b\) respectively, and any bit string that is not a pair encoding to the empty word. Then the map \(z \mapsto \langle \pi _2(z), \pi _1(z)\rangle \) is polynomial-time computable.
For all \(C, d \in \mathbb {N}\), the map sending a bit string \(x\) to the word \(1^{C(|x|+1)^d}\) of \(C(|x|+1)^d\) ones is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(K(n+1)^k\) steps on every input of length \(n\), for some constants \(K, k \in \mathbb {N}\).
For every fixed bit string \(w\), the map \(x \mapsto w\, x\) that prepends \(w\) to a bit string \(x\) is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
Let \(\sigma \) be a finite set of states, \(\delta : \sigma \times \{ 0,1\} \to \sigma \) a transition function, \(o : \sigma \times \{ 0,1\} \to \{ 0,1\} \cup \{ \bot \} \) an emission function (\(\bot \) meaning no output) and \(q_0 \in \sigma \), and let \(\tau _{\delta ,o}(q_0, x)\) be the word emitted by the one-pass transducer reading \(x \in \{ 0,1\} ^*\) from left to right starting in \(q_0\), where each bit \(b\) read in state \(p\) emits \(o(p,b)\) (or nothing) and moves to \(\delta (p,b)\). Then the map \(x \mapsto \tau _{\delta ,o}(q_0, x)\) is polynomial-time computable: some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
The map sending a bit string \(x\) to the word \(1^{|x|}\) of \(|x|\) ones is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
For languages \(L, L' \subseteq \{ 0,1\} ^{*}\), we say \(L\) is polynomial-time (Karp) reducible to \(L'\), written \(L \le _p L'\), if there is a function \(f : \{ 0,1\} ^{*} \to \{ 0,1\} ^{*}\) that is polynomial-time computable (some finite-state binary-alphabet Turing machine computes it within \(C \cdot (n+1)^c\) steps on inputs of length \(n\), for constants \(C, c \in \mathbb {N}\)) such that for every binary string \(x\), \(x \in L\) if and only if \(f(x) \in L'\) [AB09, Definition 2.7].
For \(c, C \in \mathbb {N}\), \(\mathcal{P}_{c,C}\) is the finite binary-alphabet machine with \(c+1\) work tapes, control states copy, setup, \(\mathrm{loop}_i\), \(\mathrm{rewind}_i\) (\(0 \le i \le c\)), \(\mathrm{advance}_i\) (\(0 \le i \le c+1\)) and \(\mathrm{emit}_j\) (\(0 \le j \le C\)), and initial state copy. In copy it scans the input, writing \(1\) on every work tape and moving all heads right; at the end of the input it writes one more \(1\) everywhere, moves all heads left and enters setup, which moves all work heads left until head \(0\) reads a blank, then one cell right into \(\mathrm{loop}_c\). In \(\mathrm{loop}_i\), if head \(i\) reads a blank it moves head \(i\) left into \(\mathrm{rewind}_i\), and otherwise it passes to \(\mathrm{loop}_{i-1}\), or to \(\mathrm{emit}_C\) when \(i = 0\); \(\mathrm{emit}_j\) for \(j \gt 0\) outputs a \(1\) and passes to \(\mathrm{emit}_{j-1}\), and \(\mathrm{emit}_0\) passes to \(\mathrm{advance}_0\); \(\mathrm{rewind}_i\) moves head \(i\) left until it reads a blank, then one cell right into \(\mathrm{advance}_{i+1}\); \(\mathrm{advance}_i\) for \(i \le c\) moves head \(i\) one cell right into \(\mathrm{loop}_i\), and \(\mathrm{advance}_{c+1}\) halts.
For binary strings \(w\) and \(x\), an optional state \(q \in \{ 0, 1, \ldots , |w|\} \cup \{ \bot \} \) (\(\bot \) meaning halted), an input-head position \(p \in \{ 0, 1, \ldots , |x|+1\} \) (positions \(0\) and \(|x|+1\) being the blank cells flanking \(x\)) and an output string \(u\), this is the configuration of a machine with no work tapes and state set \(\{ 0, 1, \ldots , |w|\} \) on input \(x\) that is in state \(q\), has its input head at position \(p\), and has \(u\) on its output tape.
For a binary string \(w = w_0 w_1 \cdots w_{m-1}\), the prefixing machine \(P_w\) is the binary-alphabet Turing machine with no work tapes, states \(\{ 0, 1, \ldots , m\} \) and start state \(0\). In state \(i \lt m\) it emits \(w_i\) and moves to state \(i+1\) without moving its input head; in state \(m\), if the input head scans a bit \(b\) it emits \(b\), moves the input head right and stays in state \(m\), and if it scans a blank it halts without emitting anything.
Let \(w = w_0 \cdots w_{m-1}\) be a binary string and \(P_w\) the machine with no work tapes and states \(\{ 0, \ldots , m\} \) which first emits \(w_0, \ldots , w_{m-1}\) one per step without moving its input head, then copies the input bit by bit while moving right, halting at the first blank. Then for every binary input \(x\), \(P_w\) halts within \(|w| + |x| + 1\) steps with output the concatenation \(wx\).
Let \(w = w_0 \cdots w_{m-1}\) and \(x = x_1 \cdots x_n\) be binary strings, and let \(P_w\) be the machine with no work tapes and states \(\{ 0, \ldots , m\} \) which emits \(w_i\) and advances to state \(i+1\) in each state \(i \lt m\), and in state \(m\) emits the scanned input bit and moves its input head right, halting at a blank. For every \(i \le n\), running \(P_w\) on input \(x\) for \(i\) steps from the configuration in state \(m\) with the input head on the first input cell and output \(w\) yields the configuration in state \(m\) with the input head on cell \(i+1\) and output \(w\, x_1 \cdots x_i\).
Let \(w = w_0 \cdots w_{m-1}\) and \(x\) be binary strings, and let \(P_w\) be the machine with no work tapes and states \(\{ 0, \ldots , m\} \) which, starting in state \(0\), emits \(w_i\) and moves to state \(i+1\) in each state \(i \lt m\), and in state \(m\) copies the input bit by bit, halting at the first blank. For every \(i \le m\), after \(i\) steps on input \(x\) the machine \(P_w\) is in state \(i\), its input head is still on the first input cell, and its output tape holds the first \(i\) bits \(w_0 \cdots w_{i-1}\) of \(w\).
Let \(V \subseteq \{ 0,1\} ^*\) be a language in \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time, and let \(g : \{ 0,1\} ^* \to \{ 0,1\} ^*\) be polynomial-time computable (computed by some finite binary Turing machine within \(C(n+1)^c\) steps on inputs of length \(n\), for some constants \(C, c\)). Then the preimage \(g^{-1}(V) = \{ x \in \{ 0,1\} ^* : g(x) \in V\} \) also belongs to \(\mathsf{P}\).
For a source machine \(M\), a word \(x\), an optional schedule phase, an input head position, unary and guide head positions, a unary tape, a data register, data tapes, and an output list, the preparation condition records the information preserved during initialization: except in the very first phase, the unary tape carries its origin sentinel exactly at coordinate \(0\); the data register and output are still initial; the data tapes hold blank payloads before copying, the copied payloads of the appropriate count during copying, and the full input copy in the later allocation phases; and the input, unary, and guide heads sit at their phase-prescribed positions. During the macrostep cycle only the lower bound \(1 \le u\) on the unary head is kept, and at the counter-test phase the full copy is required exactly when \(u = 1\); a halted schedule imposes no condition.
For a schedule configuration \(c\) on the embedded input, the preparation invariant is the preparation condition instantiated at the state of \(c\), its input head position, its unary and guide head positions, and its unary tape content, together with the given data register, data tapes, and output.
Applying to a schedule configuration a non-clock action of the schedule controller — with a target phase, an input move, and optional writes and moves on the three distinguished tapes — the preparation invariant of the resulting configuration equals the preparation condition at the target phase, the moved input position, the unary and guide head positions shifted by the action’s moves, and the unary tape after the action’s optional write, with the data register, data tapes, and output unchanged.
For all binary machines \(W\) and \(M\), padding constant \(a\), word \(x\), and time \(t\), there exist a data register, data tapes, and an output list such that the \(t\)-step run of the logical simulator from its initial configuration on the embedded input equals the decorated form of the schedule controller’s \(t\)-step run with these components, and the preparation invariant holds for that schedule configuration with these components.
Let a schedule configuration \(c\) on the embedded input be in a live phase and satisfy the preparation invariant with a given data register, data tapes, and output. Then the preparation invariant holds for the configuration after one step of the schedule controller, taken with the register, guide-position tape writes, and appended emission produced by the data transducer’s visit at \(c\).
Let \(M\) be a binary multi-tape Turing machine with \(k\) work tapes indexed by \(\mathbb {Z}\), let \(n, t \in \mathbb {N}\) and \(\tau \in \{ 0, \dots , k-1\} \), and for \(s \in \mathbb {N}\) let \(w(s)\) be the position of the head of work tape \(\tau \) after \(s\) steps of \(M\) on the all-zeros input \(0^n\). The last previous visit \(\mathrm{prev}_M(n, t, \tau )\) is the largest \(s \lt t\) with \(w(s) = w(t)\), and is undefined when no such \(s\) exists.
Let \(M\) be a binary multi-tape Turing machine with \(k\) work tapes indexed by \(\mathbb {Z}\), let \(n, t \in \mathbb {N}\) and \(\tau \in \{ 0, \dots , k-1\} \), and for \(s \in \mathbb {N}\) let \(w(s)\) be the position of the head of work tape \(\tau \) after \(s\) steps of \(M\) on the all-zeros input \(0^n\), and let \(\mathrm{prev}\) be the largest \(s \lt t\) with \(w(s) = w(t)\), undefined when there is no such \(s\). If \(\mathrm{prev}\) is undefined, then \(w(r) \ne w(t)\) for every \(r \lt t\).
Let \(M\) be a binary multi-tape Turing machine with \(k\) work tapes indexed by \(\mathbb {Z}\), let \(n, t \in \mathbb {N}\) and \(\tau \in \{ 0, \dots , k-1\} \), and for \(s \in \mathbb {N}\) let \(w(s)\) be the position of the head of work tape \(\tau \) after \(s\) steps of \(M\) on the all-zeros input \(0^n\), and let \(\mathrm{prev}\) be the largest \(s \lt t\) with \(w(s) = w(t)\), undefined when there is no such \(s\). If \(\mathrm{prev}\) is defined and equals \(s\), then \(s \lt t\), \(w(s) = w(t)\), and \(w(r) \ne w(t)\) for every \(r\) with \(s \lt r \lt t\).
The language \(\mathrm{SAT}\) over the binary alphabet is the set of all strings whose total CNF decoding is a satisfiable formula [AB09, Section 2.3.1]. Decoding is total, with the empty (satisfiable) formula as fallback, so every non-well-formed string belongs to \(\mathrm{SAT}\).
The language \(\mathrm{3SAT}\) over the binary alphabet is the set of all strings whose decoded CNF formula has width at most three (every clause carries at most three literal occurrences) and is satisfiable [AB09, Section 2.3.1]. The fallback formula has no clauses, so non-well-formed strings belong to \(\mathrm{3SAT}\) as well.
The language \(\mathrm{3SAT}\) — binary strings whose decoded formula is a satisfiable 3CNF — belongs to \(\mathrm{NP}\), with the same certificate parameters as \(\mathrm{SAT}\): a satisfying assignment of exactly \(|x| + 1\) bits, verified in polynomial time after an additional width check. [AB09, Theorem 2.10, membership]
A binary string \(x\) belongs to \(\mathrm{3SAT}\) if and only if there exists a certificate \(u\) with \(|u| = |x| + 1\) such that \(x \mathbin {+\! \! +} u\) belongs to the 3SAT-mode verifier language. The width condition depends only on the instance \(x\), so the same certificate as for \(\mathrm{SAT}\) suffices.
For all naturals \(c, A, d, n\) and any length \(\mathrm{len} \le A\, (n+1)^d\):
i.e. the bridge overhead is bounded by one monomial of degree one higher.
Let a multi-tape machine, started in a live configuration, be halted after \(t\) steps. Then there is a positive time \(u \le t\) such that the machine is not halted at any earlier time, and the \(u\)-step run equals the \(t\)-step run in every field: a padded halting run can be cut at its first halt without changing any endpoint.
For every binary string \(x\), the CNF formula decoded from \(x\) is satisfiable if and only if there exists a word \(u\) of length exactly \(|x| + 1\) such that the decoded formula evaluates to true under the assignment that reads variable \(v\) off the \(v\)-th bit of \(u\) (default false). [AB09, Theorem 2.10, membership]
Let \(M\) and \(U\) be finite multi-tape Turing machines over the binary alphabet, let \(f, g\) be functions on binary strings and \(T_1, T_2 : \mathbb {N} \to \mathbb {N}\) time bounds. If \(M\) computes \(f\) within \(T_1\), and for every input \(x\) the machine \(U\) run on \(f(x)\) halts within \(T_2(|x|)\) steps with output \(g(x)\), then some finite machine computes \(g\) within time \(n \mapsto 2\, T_1(n) + T_2(n) + 2\).
Let a finite multi-tape machine \(M\) compute output \(y\) on input \(x\) within \(T\) steps. Then there is a time \(t \le T\) such that \(M\) is still live at every time before \(t\), halts exactly at time \(t\), and its output at time \(t\) is \(y\): the first halting transition can be selected while retaining the final emission.
Let \(\varphi \) be a CNF formula, \(C\) a clause of \(\varphi \), and \(\ell \) a literal of \(C\). Then the variable index of \(\ell \) is strictly below the variable bound of \(\varphi \) (the maximum over all occurring indices \(v\) of \(v + 1\)).
If the fuel-indexed formula parser succeeds on a binary string \(x\), returning a formula \(\varphi \) and a remainder \(r\), then \(x\) is exactly the serialization of \(\varphi \) — every clause marker and the final terminator — followed by \(r\).
The following four functions on binary strings are polynomial-time computable: the single-bit split-validity test, the instance projection, the single-bit syntax scan of the projected instance, and the safe evaluator request assembly.
Let \(p\) be a Boolean predicate on binary strings and \(f, g\) functions on binary strings. If the single-bit function \(x \mapsto [p(x)]\) and both \(f\) and \(g\) are polynomial-time computable, then so is the function mapping \(x\) to \(f(x)\) when \(p(x)\) holds and to \(g(x)\) otherwise.
The function sending a word \(z\) to the suffix of its second pair projection obtained by dropping as many bits as the length of its first pair projection is polynomial-time computable on every word, malformed pair inputs being interpreted through the empty-default projections.
Let \(f\) be a function from binary strings to binary strings. If there exist a finite multi-tape Turing machine over the binary alphabet and a constant \(C\) such that the machine computes \(f\) within \(C\cdot (n+1)\) steps on every input of length \(n\), then \(f\) is polynomial-time computable.
If a function \(g\) on binary strings is polynomial-time computable, then so is the function that decodes its input as a pair \((a, b)\), returns the pair encoding of \(a\) with \(g(b)\), and returns the empty word when the input is not a valid pair encoding.
For a binary string \(x\), let \(\varphi _x\) be the CNF formula decoded from \(x\) and \(N(\varphi _x)\) its variable bound (one more than the largest variable index occurring in \(\varphi _x\), or \(0\) if no variable occurs). The function sending every binary string \(x\) to the unary word \(1^{N(\varphi _x)}\) is polynomial-time computable, on all strings including malformed ones.
\(\mathrm{SAT} \le _p \mathrm{3SAT}\): there is a polynomial-time computable map on binary strings — clause splitting with fresh variables — under which membership in \(\mathrm{SAT}\) of any string is equivalent to membership in \(\mathrm{3SAT}\) of its image. [AB09, Lemma 2.14]
For naturals \(A, d, D, n, m\) with \(d \le D\) and \(m + 1 \le 32\, (n+1)\):
Every polynomial on a linearly bounded request length fits the common power in the original input length.
Let \(A\) and \(B\) be multi-tape machines with the same tape count and state type, \(c\) a configuration, and \(t\) a time such that at every time \(i \lt t\) the one-step transitions of \(B\) and \(A\) agree on the configuration reached by \(A\) after \(i\) steps. Then the \(t\)-step runs of \(B\) and \(A\) from \(c\) coincide. No eventual halt is assumed.
For every native input \(x\) and existing word \(w\) on the data tape, there is a positive time \(t \le 2|w| + 3|x| + 6\) such that the appender, started at its canonical entry over the state word of \(w\), is not halted before \(t\) and at time \(t\) halts in the canonical configuration holding the state word of \(w \mathbin {+\! \! +} x\), with both the input and work heads returned to their origins and no output.
Let the native input decompose as \(x = p \mathbin {+\! \! +} r\), with the appender in its copy state at clamped input offset \(|p|\) and work tape holding \(w \mathbin {+\! \! +} p\) with the head at its right blank. Then exactly \(|r| + 1\) steps append the remaining suffix \(r\) to the work tape and enter the work-tape rewind with the head on the last occupied cell of \(w \mathbin {+\! \! +} x\).
For every \(r \le |w|\): from the rewind state with the work head at offset \(r - 1\) over a tape holding \(w\), exactly \(r + 1\) steps return the head to the origin and enter the input-rewind state; the blank at offset \(-1\) is used only locally and never written.
Let the work tape hold a word \(w = p \mathbin {+\! \! +} r\) with the head at offset \(|p|\) in the seek state. Then exactly \(|r| + 1\) steps bring the head to the first blank at offset \(|w|\) and enter the copy state, without moving the native input head.
A five-state, one-tape administrative machine that appends the native input to the word already on its work tape and then rewinds both heads: seek the first blank after the existing word, copy the input, rewind the work head over the occupied prefix, and rewind the input head. The only occupied tape is the data tape; no origin marker is introduced. It is used to form a round request and at startup of the streaming host.
Given a head literal \(h\), a list \(R\) of remaining literals, and a fresh-variable cursor \(n \in \mathbb {N}\), the clause-splitting chain is defined by recursion on \(R\) and returns a CNF formula together with the next unused cursor. If \(R\) has at least three literals \(b, c, d, \dots \), the chain is the link clause \([h,\, b,\, (n,\mathrm{true})]\) followed by the chain on head \((n,\mathrm{false})\), remaining literals \(c, d, \dots \), and cursor \(n + 1\), whose returned cursor is kept. Otherwise the chain is the single clause consisting of \(h\) followed by \(R\), with the cursor \(n\) unchanged. [AB09, Section 2.3.5, proof of Lemma 2.14]
Let a clause (head plus remaining literals) have all variable indices below \(n\), and let \(a\) be an assignment satisfying it. Then there is an assignment \(a'\) agreeing with \(a\) on all indices below \(n\) that satisfies the formula emitted by the chain construction at cursor \(n\).
Let a clause with head literal, a second literal \(b\), and a tail have all its variable indices below \(n\), and let \(a\) be an assignment satisfying it. Then there is an assignment \(a'\) agreeing with \(a\) below \(n\) such that \(a'\) satisfies both the link clause \([\mathrm{head},\, b,\, (n,\mathrm{true})]\) and the continuation clause headed by \((n,\mathrm{false})\) followed by the tail; the fresh value \(a'(n)\) is the truth of the tail under \(a\).
Let \(a\) be a truth assignment satisfying the formula emitted by the chain construction on a head literal, remaining clause, and cursor. Then \(a\) satisfies the original clause (head followed by the remaining literals). No freshness hypothesis is needed in this direction.
The uniform contract for a host module \(M\) with result and emission functions and a budget \(B\): \(M\) has a positive tape count, and for every native input and argument word there is a positive time within \(B\) of the argument length at which \(M\), started at its entry over the state word of the argument, first halts — in the canonical configuration over the state word of the result, having emitted exactly the prescribed chunk, with the native input untouched.
For every polynomial-time computable word function \(f\) there exist a module \(M\) and constants \(A, d\) such that \(M\) satisfies the clean-module contract with the identity as result function, emission \(f\), and budget \(n \mapsto A\, (n+1)^d\): the emit-mode module preserves its argument and forwards the exact computed chunk, with cleanup and the final halt inside the polynomial budget.
For every polynomial-time computable word function \(f\) there exist a module \(M\) and constants \(A, d\) such that \(M\) satisfies the clean-module contract with result function \(f\), empty emission, and budget \(n \mapsto A\, (n+1)^d\) on the argument length: every polynomial word computation has a clean first-halting install module.
Let a machine \(C\) with entry and exit states satisfy a first-positive-exit bridge contract: on every input and argument it reaches, within budget \(B\), the exit over the state word of the result with the prescribed emission, avoiding the exit at intermediate positive times. Then the call-stopping wrapper of \(C\) satisfies the uniform clean-module contract with budget \(B + 1\).
For all binary words \(u\) and \(v\), the suffix extractor run on the pair encoding \(\mathrm{pair}(u,v)\) (every bit of \(u\) doubled, then the separator \(01\), then \(v\)) halts within \(3(\ell + 1)\) steps, where \(\ell \) is the length of that paired input, and its output is the suffix of \(v\) obtained by dropping its first \(|u|\) bits.
Let the native input decompose as \(x = p \mathbin {+\! \! +} r\) with the suffix extractor in its copy state at clamped offset \(|p|\). Then exactly \(|r| + 1\) steps emit every bit of \(r\) and halt on the right blank, with the counter untouched.
Let \(x\) be an input and \(p, n \in \mathbb {N}\). From the skip state at clamped offset \(p\), with a unary counter of length \(n\) whose head is at the first blank \(n\) and empty output, one step of the suffix extractor yields the configuration in the copy state with every other field unchanged; in particular the input head does not move past the first unconsumed payload bit.
Let \(x\) be an input of the form \(p\), followed by two copies of a bit \(b\), followed by a word \(r\), and let \(n \in \mathbb {N}\). From the counting state at offset \(|p|\) with a unary counter of length \(n\), counter head at \(n\), and empty output, two steps of the suffix extractor yield the counting state at offset \(|p| + 2\) with counter length \(n + 1\) and counter head at \(n + 1\); nothing is emitted.
Let \(x\) be an input, and consider a dropper action with no counter write, input move \(m\), counter-head move \(d\), optional output bit \(b\), and target state \(q'\). If the move \(m\) takes the clamped offset \(i\) to the clamped offset \(i'\), and \(h + d = h'\), then applying the action to the canonical dropper configuration at offset \(i\) with counter length \(n\), counter head \(h\), and output \(o\) yields the canonical configuration in state \(q'\) at offset \(i'\) with the same counter, head at \(h'\), and output \(o\) followed by \(b\) when present.
Let \(u\) and \(v\) be binary words, with \(u = p \mathbin {+\! \! +} r\) split into a prefix \(p\) and a suffix \(r\), and consider the input the pair encoding \(\mathrm{pair}(u,v)\) (every bit of \(u\) doubled, then the separator \(01\), then \(v\)). Started in its counting state at input offset \(2|p|\), with a unary counter of length \(|p|\) and counter head at \(|p|\) and empty output, the suffix extractor after exactly \(2|r|\) steps is again in its counting state, at input offset \(2|u|\), with a unary counter of length \(|u|\) and counter head at \(|u|\); nothing is emitted.
Let \(u\) and \(v\) be binary words and consider the input the pair encoding \(\mathrm{pair}(u,v)\) (every bit of \(u\) doubled, then the separator \(01\), then \(v\)). Started in its counting state at input offset \(2|u|\), with a unary counter of length \(|u|\), counter head at \(|u|\), and empty output, the suffix extractor after exactly two steps is in its counter-rewind state at input offset \(2|u| + 2\) (the first payload bit, just past the separator), with the counter unchanged, counter head at \(|u| - 1\), and nothing emitted.
For every \(r \le n\): in the skip state with counter length \(n\), \(r\) steps advance the clamped payload offset by \(r\) while walking the counter head to \(r\). Native clamping makes this total even when the stored length exceeds the payload.
Let \(x\) be an input, \(i\) a natural offset, and \(n \in \mathbb {N}\). Applying the dropper action that writes a one under the counter head, moves the input head and the counter head right, and enters the counting state, to a canonical dropper configuration on \(x\) at clamped offset \(i\) whose unary counter has length \(n\) and head at \(n\) (with empty output), yields the canonical configuration in the counting state at clamped offset \(i + 1\) with counter length \(n + 1\) and counter head at \(n + 1\): writing the counter’s right blank appends exactly one unary cell.
Given an input-head move \(m\), a counter-head move \(d\), an optional counter write, an optional output bit, and an optional next state among six control states, the dropper action is the one-tape action performing exactly these operations; these are the actions of the length-controlled suffix extractor.
For an input \(x\), an optional control state \(q\), natural numbers \(i\) and \(n\), an offset \(h \in \mathbb {Z}\), and an output word \(o\), the canonical dropper configuration is the configuration of the suffix extractor in control \(q\), with the input head at the clamped offset \(\min (i, |x|)\), the unary counter tape of length \(n\) with head at \(h\), and emitted output \(o\).
The suffix-extraction machine is a six-state finite machine with one work tape. Given binary words \(u\) and \(v\) and the input the pair encoding \(\mathrm{pair}(u,v)\) (every bit of \(u\) doubled, then the separator \(01\), then \(v\)), it counts the doubled prefix on a unary counter tape, recognizes the separator, rewinds the counter, skips exactly \(|u|\) payload positions (clamping at the right boundary), and copies the remaining suffix of \(v\). No output precedes recognition of the separator.
Let \(u\) be the certificate word on the evaluator’s last tape and \(C\) a clause all of whose variable indices are below \(|u|\), with serialization \(\mathrm{ser}(C)\). Suppose the evaluator’s input is a word \(p\), followed by the doubling of \(\mathrm{ser}(C)\), followed by a word \(r\). Started at input position \(|p|\) in the clause phase with formula bit \(a\), clause bit \(c\), and certificate head at \(0\), there is a number \(t \le 3\, |\mathrm{ser}(C)|\) of steps after which the evaluator is in the formula phase with formula bit \(a \wedge (c \vee C(u))\), where \(C(u)\) is the truth value of \(C\) under the assignment read off \(u\), with the clause bit reset, input position \(|p| + 2\, |\mathrm{ser}(C)|\), and certificate head at \(0\). No output bit is emitted by this pass, even for an empty or false clause.
There exist a single finite machine \(E\) and a constant \(A\) such that for every CNF formula \(\varphi \) and every certificate word \(u\) covering all variables of \(\varphi \) (each literal index below \(|u|\)), the machine \(E\), run on the pair encoding of the serialization of \(\varphi \) with \(u\), halts within \(A \cdot (\ell + 1)\) steps — \(\ell \) the length of that paired input — and outputs the single bit giving the truth value of \(\varphi \) under the assignment read off \(u\).
Suppose the evaluator’s transition at a semantic state on the current input bit is the administrative action moving right, displacing the certificate head by \(d\), and entering a new semantic state with the skip flag set. Then two machine steps — the semantic transition followed by its silent skip — consume the doubled input bit: the input position advances by two, the certificate head by \(d\), and the control reaches the new semantic state with the skip flag cleared. The certificate head moves only on the first of the two physical transitions.
Let \(u\) be the certificate word on the evaluator’s last tape and \(\varphi \) a CNF formula all of whose variable indices are below \(|u|\), with serialization \(\mathrm{ser}(\varphi )\). Suppose the evaluator’s input is a word \(p\), followed by the doubling of \(\mathrm{ser}(\varphi )\), followed by a word \(r\). Started at input position \(|p|\) in the formula phase with formula bit \(a\), clause bit false, and certificate head at \(0\), there is a number \(t \le 3\, |\mathrm{ser}(\varphi )|\) of steps after which the evaluator is halted and its output is the single bit \(a \wedge \varphi (u)\), where \(\varphi (u)\) is the truth value of \(\varphi \) under the assignment read off \(u\); only the final formula terminator emits.
Suppose the evaluator’s input contains, at the current position, the doubling of a unary run \(1^n\). In the index phase, exactly \(2n\) steps consume the doubled run, advancing the input position by \(2n\) and the certificate head by \(n\), leaving the phase and both accumulator bits unchanged. (The first one of a literal is consumed by the clause phase, so this scan counts the variable index itself.)
Suppose the evaluator’s input contains, at the current position, the doubling of a serialized literal \((v, b)\) with \(v \lt |u|\) for the certificate \(u\). Starting in the clause phase with clause bit \(c\), exactly \(3v + 8\) steps process the literal: the input advances by \(2(v + 3)\), the certificate head walks to position \(v\), the polarity is compared with the \(v\)-th certificate bit, the clause bit becomes \(c \vee (u_v = b)\), and the certificate head returns to \(0\).
Let \(q\) be an evaluator control state whose transition moves the certificate head left while its cell is occupied and otherwise steps right into a destination state. For every \(n \le |u|\), starting from the canonical configuration with certificate head at offset \(n - 1\), exactly \(n + 1\) steps return the head to offset \(0\) in the destination state, preserving all native input and output fields.
Let \(w\) be an input and \(i \lt |w|\). From a canonical evaluation configuration at input position \(i\) in a semantic state with phase \(q\), accumulator bits \(a\) and \(c\), and the skip flag set, one evaluator step yields the canonical configuration at position \(i + 1\) in the same phase with the same accumulator bits and the skip flag cleared; the certificate head, the saved bank, and the empty output are unchanged.
Let the extractor machine \(M\) compute the certificate \(u\) on input \(w\) within \(T\) steps. Then within \(T + |w| + |u| + 5\) steps from its initial configuration, the evaluator built over \(M\) reaches the canonical evaluation configuration: some saved bank of \(M\), the certificate \(u\) on the last tape with head at \(0\), the input head at \(0\), formula phase with formula bit true, and empty output.
Let \(M\) be a base machine with \(k\) work tapes. Given an input move \(m\), a certificate-head move \(d\), an optional output bit \(b\), and an optional target state \(q\), the administrative evaluation action is the action of the evaluator over \(M\) that writes nothing, keeps the first \(k\) work heads stationary, moves the input head by \(m\) and the last (certificate) head by \(d\), emits \(b\) when present, and switches to \(q\) (halting when \(q\) is absent).
Applying a non-emitting administrative evaluator action to a canonical evaluation configuration yields the canonical configuration with the prescribed new control state, the input position moved as stated, and the certificate head displaced by the stated amount; the saved extractor bank and the empty output are preserved.
The canonical configuration of the evaluator during its streaming phase: the first \(k\) work tapes hold a saved configuration of the extractor machine, the last tape holds the immutable certificate word \(u\) with its head at a stated integer offset, the control is a stated semantic state, the input head sits before a stated position, and the output is empty.
Let \(u\) be a certificate word and \(j \in \mathbb {Z}\). In every canonical evaluation configuration with certificate \(u\) and certificate head at offset \(j\), the symbol read by the last work head is the cell of the buffer tape of \(u\) at offset \(j\), that is, the \(j\)-th bit of \(u\) when \(0 \le j \lt |u|\) and the blank otherwise.
The control-state type of the streaming formula evaluator: either one of four administrative states (capture return, input rewind, dispatch, certificate rewind), or a semantic phase — one of five streaming positions (formula, clause, unary index, polarity, rewind) — together with a formula truth bit, a clause truth bit, and a flag indicating that the second copy of a doubled input bit is to be skipped.
Given a certificate-extractor machine \(M\) with \(k\) work tapes, the evaluator is a finite machine with \(k + 1\) work tapes that first simulates \(M\), capturing its output (the certificate) on the fresh last tape, then rewinds the native input and the certificate buffer, and finally streams a previously validated doubled serialized formula, evaluating it clause by clause against the certificate: unary index walks advance the certificate head, the polarity bit is compared with the certificate cell, and clause and formula truth bits are accumulated in finite control. No physical output occurs before the single final verdict bit. [AB09, Theorem 2.10, membership]
Let the \(i\)-th module (with positive tape count), started at its own entry over the state word of an argument, stay live before time \(t\) and halt at time \(t\) over the state word of a result with output \(o\). Then the host, started at that module’s entry over the same argument with output prefix \(p\): (i) never visits the loop anchor before time \(t\), and (ii) at time \(t\) reaches the module’s return state over the state word of the result with output \(p \mathbin {+\! \! +} o\).
Let the \(i\)-th module of the host satisfy the clean-module contract with result \(f\), emission \(g\), and budget \(B\). Then for every native input, argument, and output prefix, there is a positive time within \(B\) of the argument length during which the host, started at that module’s entry, avoids the loop anchor and then reaches the module’s return state over the state word of \(f\) of the argument, with \(g\) of the argument appended to the output.
For a module with positive tape count: forwarding a padded canonical configuration of the module (built from a state word, with a given output) through the emitting embedding produces the host’s canonical seam over the same state word, with the supplied prefix prepended to the output; a halted module state maps to the module’s return state.
The finite controller hosting the six modules on one shared padded work bank: the anchor dispatches into the recurring call sequence, and inside a module every transition (including the halting one) is forwarded by the output-emitting embedding, redirecting halts to the module’s return state. The clean-call contracts restore the shared scratch before each next call.
Let the input decompose as \(x = p \mathbin {+\! \! +} r\). From any state \(q\) at position \(|p|\) with emitted prefix \(o\), the word-transducer machine halts after exactly \(|r| + 1\) steps with output \(o\) followed by the transducer semantics applied to \(q\) and \(r\).
For an input word \(x\), a control state \(q\), a position \(i \le |x|\), and an output word \(o\), the transducer configuration is the configuration of a zero-work-tape machine on input \(x\) with control state \(q\), input head just before position \(i\), and already-emitted output \(o\).
Given a finite state set \(S\) with a transition \(\delta \), a per-bit optional emission, an optional final emission, and a start state, the word-transducer machine is the finite Turing machine with no work tapes that, on each input bit, applies \(\delta \), emits the per-bit optional output, and moves right, and at the right boundary emits the optional final bit and halts.
Given a state set \(S\), a transition \(\delta \), a per-bit optional emission \(\varepsilon \), and an optional final emission \(\phi \), the transducer semantics is the word function defined by recursion on the input: on the empty word from state \(q\) it returns the final emission \(\phi (q)\) (zero or one bit), and on a bit \(b\) followed by a word \(r\) from state \(q\) it returns the emission \(\varepsilon (q, b)\) followed by the semantics of \(r\) from state \(\delta (q, b)\).
Let \(\varphi \) be a CNF formula with serialization \(\mathrm{ser}(\varphi )\) and its variable bound \(N(\varphi )\) (one more than the largest variable index occurring in \(\varphi \), or \(0\) if no variable occurs). The intercepted maximum-pass machine, run on \(\mathrm{ser}(\varphi )\), halts within \(4\, |\mathrm{ser}(\varphi )| + 6\) steps, and its output is the unary word \(1^{N(\varphi )}\).
Let \(x\) be an input, \(n \in \mathbb {N}\), and \(B\) any buffer tape. Started in the streaming entry state with the input head at \(0\), a unary counter of length \(n\) with head at \(0\), the buffer \(B\) with head at \(0\), and empty output, the intercepted maximum-pass machine after exactly \(n + 1\) steps is halted with output \(1^n\) and counter head at \(n\), all other fields unchanged; in particular the first blank after the counter halts the machine without an extra output bit.
For every \(r \le n\): from the streaming entry frame with counter length \(n\), exactly \(r\) transitions of the intercepted machine copy the first \(r\) unary counter cells to the output, moving only the counter head; no native input or buffer movement occurs during the replay.
Let \(\varphi \) be a CNF formula with serialization \(\mathrm{ser}(\varphi )\) and its variable bound \(N(\varphi )\) (one more than the largest variable index occurring in \(\varphi \), or \(0\) if no variable occurs). There is a time \(t \le 3\, |\mathrm{ser}(\varphi )| + 5\) at which the intercepted maximum-pass machine, started on \(\mathrm{ser}(\varphi )\), is in its streaming entry state with the counter tape holding \(1^{N(\varphi )}\), the buffer empty apart from its permanent marker, the input head and both work heads at \(0\), and empty output. This is the same endpoint as for the unmodified clause-splitting transducer.
The machine obtained from the clause-splitting transducer by intercepting its streaming entry state: instead of streaming, it copies the unary counter to the output (one bit per occupied cell) and halts at the first blank. All other transitions are taken verbatim from the transducer, so the verified initialization and maximum pass apply unchanged.
For every machine \(M\), tape bound \(K \ge k\), configuration \(c\), and time \(t\): running the padded machine for \(t\) steps from the padded configuration equals padding the \(t\)-step run of \(M\) from \(c\) — the padded source observes exactly its original work symbols.
For a positive source tape count, padding the canonical configuration built from a state word (data word on tape zero, other tapes blank) yields the canonical configuration over the larger tape count built from the same state word, with the same output.
Let the machine \(M\) compute a function \(f\) within time \(T\). Then the pairing machine over \(M\) computes the function \(x \mapsto \mathrm{pair}(f(x), x)\) within time \(n \mapsto 4\, T(n) + 2n + 8\): retaining a computed result next to its original input has a uniform overhead, the capture output length being paid by the source time.
Let the native input decompose as \(x = p \mathbin {+\! \! +} r\) with the pairing machine in its native-copy state at input position \(|p|\). Then exactly \(|r| + 1\) steps emit every bit of \(r\) and halt on the right blank, leaving all work tapes and the capture head unchanged.
Let the captured word decompose as \(u = p \mathbin {+\! \! +} r\) with the capture head at offset \(|p|\) in the replay state. Then exactly \(2\, |r| + 2\) steps emit the doubling of \(r\) followed by the separator \(01\), move the capture head to \(|u|\), and enter the native-copy state; the source bank and the native input head remain fixed.
For every \(n \le |u|\): from the canonical pairing configuration in the capture-rewind state with the capture head at offset \(n - 1\) and empty output, exactly \(n + 1\) steps return the head to the origin and enter the replay state, with all other fields unchanged.
Let the source machine \(M\) compute the word \(u\) on input \(x\) within \(T\) steps. Then within \(T + |x| + |u| + 5\) steps from its initial configuration, the pairing machine over \(M\) reaches the canonical pairing configuration at the replay entry: some saved bank of \(M\), the captured word \(u\) with head at the origin, the input head at the origin, and all startup output suppressed.
Let \(M\) be a base machine with \(k\) work tapes. Given an input move \(m\), a capture-head move \(d\), an optional output bit \(b\), and an optional target state \(q\), the pairing administrative action is the action of the pairing machine over \(M\) that writes nothing, keeps the completed source bank (the first \(k\) work heads) stationary, moves the input head by \(m\) and the last (capture) head by \(d\), emits \(b\) when present, and switches to \(q\).
Applying a pairing administrative action to a canonical pairing configuration changes only the control state, the two displayed head positions, and the output (appending the optionally emitted bit); the saved source bank and the captured word are preserved.
Let \(M\) be a finite machine with \(k\) work tapes and \(x\) an input. Given a configuration \(s\) of \(M\) on \(x\) (the saved source bank), a captured word \(u\), an optional control state \(q\), a position \(i \le |x|\), an offset \(j \in \mathbb {Z}\), and an output word \(o\), the canonical pairing configuration is the configuration of the pairing machine over \(M\) whose first \(k\) work tapes and heads are those of \(s\), whose last tape holds \(u\) with head at \(j\), whose control is \(q\) (halted when \(q\) is absent), whose input head is just before position \(i\), and whose output is \(o\).
Let \(M\) be a finite machine, \(x\) an input, and \(i \le |x|\). In every canonical pairing configuration over \(M\) on input \(x\) at input position \(i\), the symbol read by the input head is the \(i\)-th bit of \(x\) when \(i \lt |x|\), and the blank when \(i = |x|\).
Let \(M\) be a finite machine, \(u\) a captured word, and \(j \in \mathbb {Z}\). In every canonical pairing configuration over \(M\) with captured word \(u\) and capture head at offset \(j\), the symbol under the last work head is the cell of the buffer tape of \(u\) at \(j\), that is, the \(j\)-th bit of \(u\) when \(0 \le j \lt |u|\) and the blank otherwise.
Given a machine \(M\) computing a function \(f\), the pairing machine with one extra tape: it runs \(M\), capturing its output onto the fresh last tape, rewinds the native input and the captured word, replays each captured bit twice, emits the pairing separator \(01\), and finally copies the native input. On input \(x\) its output is therefore the pair encoding of \(f(x)\) with \(x\).
Let \(q\) be a transducer state whose transition moves the counter head left while its cell is occupied and otherwise steps right into a destination state. For every \(r \le n\), starting from the canonical frame with counter length \(n\) and counter head at offset \(r - 1\), exactly \(r + 1\) silent transitions reach the destination state with the counter head at \(0\); input position, buffer, and output are untouched.
For every input \(x\), the first two steps of the clause-splitting transducer from its initial configuration on \(x\) are silent and yield the canonical frame in the maximum-pass entry state with input head at \(0\), empty counter, the empty marked buffer (blank except for the permanent marker at offset \(-1\)), both work heads at \(0\), and empty output.
Let \(C\) be a clause with serialization \(\mathrm{ser}(C)\), and let \(m_C\) be the maximum of \(v + 1\) over the literals \((v, b)\) of \(C\) (or \(0\) if \(C\) is empty). Suppose the input is \(x = p \mathbin {+\! \! +} \mathrm{ser}(C) \mathbin {+\! \! +} r\), and the transducer is in its clause-level maximum state at input position \(|p|\), with counter length \(n\), counter head at \(0\), and arbitrary buffer, buffer head, and output. Then there is a number \(t \le 2\, |\mathrm{ser}(C)|\) of silent steps after which the transducer is in formula-level maximum control at input position \(|p| + |\mathrm{ser}(C)|\), with counter length \(\max (n, m_C)\), counter head back at \(0\), and buffer, buffer head, and output unchanged.
Let \(\varphi \) be a CNF formula with serialization \(\mathrm{ser}(\varphi )\) and its variable bound \(N(\varphi )\) (one more than the largest variable index occurring in \(\varphi \), or \(0\) if no variable occurs). Suppose the input is \(x = p \mathbin {+\! \! +} \mathrm{ser}(\varphi ) \mathbin {+\! \! +} r\), and the transducer is in formula-level maximum control at input position \(|p|\), with counter length \(n\), counter head at \(0\), and arbitrary buffer, buffer head, and output. Then there is a number \(t \le 2\, |\mathrm{ser}(\varphi )|\) of silent steps after which the transducer is at its rewind seam at input position \(|p| + |\mathrm{ser}(\varphi )| - 1\), with counter length \(\max (n, N(\varphi ))\), counter head at \(0\), and buffer, buffer head, and output unchanged.
Suppose the input contains a serialized literal \((v,\mathrm{pol})\) at the current position and the transducer is in its clause-level maximum state with counter length \(n\) and counter head at \(0\). Then exactly \(2v + 5\) transitions consume the literal, record the counter length \(\max (n,\, v + 1)\), restore the counter head to \(0\), and return to the same state with the input advanced by \(v + 3\).
Suppose the input contains a unary run \(1^v\) at the current position and the transducer is in its index-scanning state with counter length \(n\) and counter head at \(j \le n\). Then exactly \(v\) steps consume the run, advancing the input by \(v\), moving the counter head to \(j + v\), and extending the counter length to \(\max (n,\, j + v)\).
If the transducer’s transition at the current state, input symbol, and work symbols is a silent action with no writes, then one machine step from the canonical reduction frame changes only the control state and the three head positions, as prescribed; the counter length, buffer, and output are unchanged.
Let \(\varphi \) be a CNF formula with serialization \(\mathrm{ser}(\varphi )\) and its variable bound \(N(\varphi )\) (one more than the largest variable index occurring in \(\varphi \), or \(0\) if no variable occurs). There is a time \(t \le 3\, |\mathrm{ser}(\varphi )| + 5\) at which the clause-splitting transducer, started on \(\mathrm{ser}(\varphi )\), is in its streaming entry state with exactly \(1^{N(\varphi )}\) on the counter tape, the buffer empty apart from its permanent marker, the input head back at \(0\), both work heads at \(0\), and empty output.
If the clause-splitting transducer, run for \(t\) steps from its initial configuration, has empty output, then at no earlier time was its control in the streaming entry state: that state necessarily emits a bit, which output-prefix monotonicity would make visible at time \(t\).
Given an input move \(m\), optional writes \(w_0, w_1\) and head moves \(d_0, d_1\) for the two work tapes, an optional output bit \(b\), and an optional next state \(q\), the two-tape reduction action is the action of the 35-state clause-splitting transducer performing exactly these operations; work tape \(0\) is a unary fresh-variable counter and work tape \(1\) a temporary literal buffer.
Applying a two-tape reduction action to a canonical reduction frame yields the canonical frame whose input position, counter tape, buffer tape, both head offsets, and appended output are exactly as prescribed by the action’s components; this frame-level calculus is shared by the maximum pass and the streaming serializer.
The literal-buffer tape content determined by a word and a cut position: cells before the cut have been erased, cells from the cut onwards carry the word’s remaining bits, and the cell at offset \(-1\) carries a permanent marker allowing return even after the payload has been completely erased.
Given an input \(x\), an optional control state \(q\), a position \(i \le |x|\), a counter length \(n\), a buffer tape \(B\), head offsets \(a, b \in \mathbb {Z}\), and an output word \(o\), the canonical reduction frame is the configuration of the 35-state transducer in control \(q\), with the input head just before position \(i\), the unary counter tape of length \(n\) on work tape \(0\) with head at \(a\), the buffer tape \(B\) on work tape \(1\) with head at \(b\), and accumulated output \(o\).
For every \(n \in \mathbb {N}\), the canonical unary counter tape of length \(n\) is the tape content holding a one at each offset \(0, \dots , n - 1\) and a blank at every other integer offset; its occupied length \(n\) represents the next unused variable index.
Let \(n, j \in \mathbb {N}\) with \(j \le n\). Writing a one at offset \(j\) of the unary counter tape of length \(n\) yields the unary counter tape of length \(\max (n, j + 1)\): writing within the current prefix preserves it, while writing its right blank extends the length by one.
A 35-state, two-tape candidate transducer for the clause-splitting reduction on previously validated serialized CNF words: states 0–1 install a permanent left marker on the literal buffer, states 2–6 silently compute the maximum unary literal length on the counter tape, states 7–8 rewind the native input, and states 9–34 form a proposed streaming serializer. Only the initialization, the maximum pass, and the rewind are verified for this machine; it is not itself a polynomial-time witness.
The string-level SAT-to-3SAT reduction is the map on binary strings that decodes its input to a CNF formula (with the empty formula as fallback for malformed inputs), applies the clause-splitting transform (fresh variables allocated from the decoded formula’s variable bound upwards), and serializes the result. [AB09, Lemma 2.14]
The string-level SAT-to-3SAT reduction — decode, apply the clause-splitting transform, and serialize, with the empty-formula fallback on malformed inputs — is polynomial-time computable: a finite host of six clean modules (startup append and install, then the recurring pack, append, emit, and install calls) drives the normalized streaming schedule through an emitting loop with round budget \(R(n) = n\), every request being linearly bounded in the original input.
Let \(x\) be a binary word and \(s\) a seam state with fresh cursor \(f\), consumed length \(c\), and phase \(q\), and let \(w_s\) be the encoded seam word of \(s\). On the canonical round request \(\mathrm{pair}(w_s, x)\), the fresh, consumed, phase, and unread-input projections return exactly \(1^f\), \(1^c\), \(1^q\), and the word \(x\) with its first \(c\) bits dropped.
On every canonical round request — the pair encoding of the encoded seam word of any state \(s\) (not only reachable ones) with the input \(x\) — the word-level next-state program returns the encoded word of the state after one streaming round, and the word-level chunk program returns exactly the chunk emitted by that round.
A total straight-line word program computing the chunk emitted by one streaming round from a round request: finished phase emits nothing; exhausted input or a terminator bit emits \(0\); a clause marker at formula phase emits \(1\); a missing polarity bit emits nothing; a tail link emits the assembled fragment; otherwise the current literal is copied. Only the grammar phase and the complete-literal lookahead are tested locally — full validation belongs to startup.
The one larger output chunk of a tail-link round, assembled from the request fields: the positive fresh link (\(1\) followed by the unary fresh field, then \(01\)), two clause markers, the negative fresh link, and the consumed original literal, in that order.
A round request is the pair encoding of an encoded seam word with a fresh copy of the native input. For a word \(z\), the fresh field of \(z\) is the first component of the first component of \(z\) (each projection defaulting to the empty word on malformed inputs); on a canonical request it is the unary fresh cursor.
For a word \(z\), the phase field of \(z\) is the second component of the second component of the first component of \(z\) (each projection defaulting to the empty word on malformed inputs); on a canonical round request it is the unary phase tag.
A total straight-line word program installing the next encoded seam word from a round request, by the same local tests as the chunk program: the consumed field advances by the whole current literal length (or by one for markers and terminators), the phase tag is updated, and the fresh field grows by one exactly on a tail link.
For a word \(z\), the consumed field of \(z\) is the first component of the second component of the first component of \(z\) (each projection defaulting to the empty word on malformed inputs); on a canonical round request it is the unary consumed length.
The safe input handed to the evaluator: the catalog split of the string when the full validity guard holds, and otherwise the pair encoding of the serialized empty formula with the empty certificate. Thus malformed formulas are replaced before the evaluator runs, and no semantic rejection can precede validation.
For every binary string \(z\) there exist a CNF formula \(\varphi \) and a word \(u\) such that the safe evaluator request on \(z\) is exactly the pair encoding of the serialization of \(\varphi \) with \(u\), every variable index occurring in \(\varphi \) is below \(|u|\), and the truth value of \(\varphi \) under the assignment read off \(u\) equals the safe evaluation value of \(z\). This holds also on failed splits and failed parses.
The Boolean value of the evaluation after the syntax pass: when the full validity guard holds, the truth value of the decoded instance under the assignment read off the recovered certificate; otherwise true (the fallback empty formula evaluates to true).
The single-bit function sending a binary string \(z\) to its safe evaluation value is polynomial-time computable: the uniform evaluator is applied only on safe requests, which are certified serializations with covering certificates.
Let \(S\) be a finite set with transition function \(\delta : S \times \{ 0,1\} \to S\), start state \(s_0 \in S\), and accepting predicate \(\alpha \) on \(S\). On every input \(x\) of length \(n\), the one-way scanner built from these data halts within \(n + 1\) steps, and its output is the single bit \(\alpha (\delta ^*(s_0, x))\), where \(\delta ^*(s_0, x)\) is the state reached by folding \(\delta \) over \(x\) from \(s_0\).
Let the input decompose as \(x = p \mathbin {+\! \! +} r\). Started in any state \(q\) at position \(|p|\), the one-way scanner halts after exactly \(|r| + 1\) steps with output the single bit obtained by applying the accepting predicate to the fold of the transition over \(r\) from \(q\); no earlier step emits.
For an input word \(x\), a control state \(q\), and a position \(i \le |x|\), the scanner configuration is the configuration of a zero-work-tape machine on input \(x\) in control state \(q\), with the input head just before position \(i\) and empty output.
Given a finite state set \(S\) with a transition \(S \times \{ 0,1\} \to S\), a start state, and an accepting predicate on \(S\), the one-way scanner is the finite Turing machine with no work tapes that reads its input left to right applying the transition, and upon reaching the right boundary emits the single bit given by the accepting predicate at the final state and halts.
Let a machine \(C\), started at its entry over the state word of an argument, avoid its designated exit at all intermediate positive times before \(t \gt 0\) and reach at time \(t\) the exit over the state word of a result with a given output. Then the call-stopping wrapper, started with the flag down, runs for \(t + 1\) steps and halts in that same physical configuration (flag raised, state halted), retaining every emitted bit.
Given a machine \(C\) with designated entry and exit states, the wrapper machine whose states carry an additional Boolean release flag: it simulates \(C\), raising the flag after the first step, and halts silently as soon as the flag is raised and the control is at the designated exit. The flag ensures that even an entry equal to the exit executes at least once.
For every binary string \(x\), concatenating over \(i \lt |x| + 1\) the chunks produced by the encoded chunk function at the \(i\)-fold iterate of the encoded next-state function on the encoded startup word equals the image of \(x\) under the string-level SAT-to-3SAT reduction — exactly the expression returned by the emitting-loop constructor with round budget \(R(n) = n\).
For a clause \(C\) and fresh cursor \(j\): the part of the transformed clause’s serialization emitted after the initial clause marker (which is handled by the formula phase), including its final terminator, together with the final cursor. Up to two leading literals are copied verbatim before the pending tail output takes over.
For every clause \(C\) and cursor \(j\): concatenating a clause marker with the serialization of each clause of the split of \(C\) at \(j\) yields exactly one marker followed by the emitted clause body of \(C\) at \(j\), and the split’s returned cursor equals the body’s final cursor. This includes the empty clause and all clauses of width at most three.
For a CNF formula \(\varphi \) with clauses \(C_1, \dots , C_m\), the round count of \(\varphi \) is \(1 + \sum _{k=1}^{m} (|C_k| + 2)\): per clause one marker round, one round per literal occurrence, and one terminator round, plus one final formula-terminator round.
The total decoding of a word into a seam state: the lengths of the two unary counter fields give the fresh cursor and consumed length, and the length of the phase field is reduced modulo five. The modulus only totalizes malformed words and is the identity on all reachable tags.
One round of the streaming clause-splitting schedule on an input string and a seam state: the finished phase is absorbing with empty output; an exhausted input emits the fallback terminator \(0\) and finishes; a terminator bit \(0\) is copied, closing a clause or the whole formula; a clause marker \(1\) at formula phase is copied; and otherwise a complete literal is parsed together with a one-bit lookahead — in the tail phase with more literals pending, the round allocates one fresh variable and emits the positive link, two clause markers, the fresh negative literal, and the consumed original literal, while in all other literal positions the literal is copied verbatim. Local parse failures terminate silently.
Let the input be \(p \mathbin {+\! \! +} 0 :: r\) and the seam state have consumed \(|p|\) bits in a non-finished phase \(q\). The round emits \([0]\), consumes one bit, and moves to the finished phase when \(q\) was the formula phase and back to the formula phase otherwise.
Let the input be \(p\) followed by a serialized literal \(\ell \) and a remainder \(r\), and let the seam state have consumed \(|p|\) bits in a literal phase (neither formula nor finished). If the phase is the tail phase and \(r\) starts with a clause bit \(1\), the round allocates the fresh cursor \(j\): it emits the positive link \((j,\mathrm{true})\), two markers, the negative literal \((j,\mathrm{false})\), and \(\ell \), incrementing the cursor; otherwise it emits \(\ell \) verbatim and advances the phase. In both cases the consumed length advances by the serialized literal length, so no literal buffer survives a round boundary.
Let the input contain a serialized clause \(C\) at the consumed position, with the seam in the first-literal phase at cursor \(j\). Then exactly \(|C| + 1\) rounds consume the serialized clause, emit exactly the emitted clause body of \(C\) at \(j\), return to the formula phase, and leave the cursor at the body’s final value.
For every binary string \(x\), running \(|x| + 1\) streaming rounds from the startup seam state of \(x\) emits exactly the image of \(x\) under the string-level SAT-to-3SAT reduction. This is an exact output identity on every string, prior to any machine-computability claim.
Let the input contain the serialization of a formula \(\varphi \) at the consumed position, with the seam in the formula phase at cursor \(j\). Then exactly the round count of \(\varphi \) many rounds consume the serialized formula, emit exactly the serialization of the threaded clause transform of \(\varphi \) at \(j\), reach the finished phase, and leave the cursor at the transform’s final value.
For every input \(x\), round count \(k\), and seam state \(s\), the output of \(k\) streaming rounds equals the concatenation over \(i \lt k\) of the chunk emitted by the round applied to the \(i\)-fold iterate of the next-state map on \(s\) — the range-indexed form required by the emitting-loop contract.
Let \(\varphi \) be a CNF formula with serialization \(\mathrm{ser}(\varphi )\) and its variable bound \(N(\varphi )\) (one more than the largest variable index occurring in \(\varphi \), or \(0\) if no variable occurs). Running \(|\mathrm{ser}(\varphi )| + 1\) streaming rounds on input \(\mathrm{ser}(\varphi )\), starting from the formula phase with fresh cursor \(N(\varphi )\) and nothing consumed, emits a word equal to the serialization of the clause-splitting transform of \(\varphi \): the semantic schedule finishes within the budget and all remaining rounds are empty.
Let the input contain a serialized clause tail \(C\) at the consumed position, with the seam in the tail phase at cursor \(j\). Then exactly \(|C| + 1\) rounds — one per remaining literal plus one for the clause terminator — consume the serialized tail, emit exactly the pending tail output of \(C\) at \(j\), return to the formula phase, and leave the cursor at the tail’s final value.
The seam state selected before any emission: if the complete syntax scan of the input accepts, start in the formula phase with the fresh cursor at the decoded formula’s variable bound and nothing consumed; otherwise start in the formula phase at the right input boundary, whose first round emits only the fallback terminator and then finishes — even at input length zero.
A streaming seam state is a triple consisting of a fresh-variable cursor in \(\mathbb {N}\), a number of consumed input bits in \(\mathbb {N}\), and one of five grammar phases (formula, first literal, second literal, tail, finished); it is the persistent state carried across one round of the normalized streaming schedule.
For a remaining clause tail and a fresh cursor \(j\), the pending tail output is the output still to be emitted after the first two literals of a clause, together with the final cursor. An empty tail emits just the clause terminator \(0\); a single literal is copied and followed by the terminator; otherwise one link is closed with the fresh positive literal \((j,\mathrm{true})\), a new clause is opened with \((j,\mathrm{false})\), the next original literal is copied, and the recursion continues on the rest of the tail at cursor \(j + 1\). [AB09, Section 2.3.5, proof of Lemma 2.14]
For literals \(a, b\), a clause tail \(C\), and a cursor \(j\): concatenating a clause marker with the serialization of each clause of the chain construction on \(a\), \(b :: C\), \(j\) yields exactly one marker followed by the serializations of \(a\) and \(b\) and the pending tail output of \(C\) at \(j\); moreover the chain’s returned cursor equals the tail’s final cursor.
The single-bit function returning the Boolean result of the complete CNF syntax scan is polynomial-time computable; the witnessing machine is a one-way finite-state scanner, and the result concerns every input string, not merely serialized formulas.
For every binary string \(x\), the Boolean syntax scan of \(x\) returns true exactly when the exact-consumption CNF parser succeeds on \(x\). This holds on every string, including the empty input, unfinished literals, and strings with trailing garbage.
The transition function of the six-state CNF syntax automaton over \(\{ 0,1\} \) is defined on the positions formula, clause, unary index, polarity, exact end, and error as follows. At formula level a \(1\) leads to clause level and a \(0\) to exact end; at clause level a \(1\) leads to the index position and a \(0\) back to formula level; the index position stays on \(1\) and moves to polarity on \(0\); polarity returns to clause level on either bit; exact end and error move to error on every bit.
For each of the six scanner positions, the residual grammar of that position is the set of binary strings that may legally remain to be read: at formula level, the serializations of CNF formulas; at clause level, a serialized clause followed by a serialized formula; at index level, a unary run \(1^v\) followed by \(0\), a polarity bit, a serialized clause, and a serialized formula; at polarity level, one arbitrary bit followed by a serialized clause and a serialized formula; at exact end, only the empty string; at the error position, no string at all.
If the literal parser fails on a string starting with a one (a literal marker), then the unary token splitter on that string returns an empty suffix: there is no polarity bit, which keeps the total round implementation faithful even on locally malformed state–input combinations.
For every literal \(\ell = (v, b)\) and remainder \(r\), the unary token splitter applied to the serialization of \(\ell \) followed by \(r\) returns the token \(1^{v+1}0\) together with the suffix \(b :: r\): the delimiter is consumed and the polarity bit remains the first bit of the returned suffix.
For a CNF formula \(\varphi \) and a fresh-variable cursor \(n\), the threaded clause transform of \(\varphi \) at \(n\) is obtained by splitting each clause of \(\varphi \) in order, threading one global fresh-variable cursor from clause to clause starting at \(n\), and concatenating the resulting 3CNF pieces; the final cursor is returned alongside. [AB09, Section 2.3.5, proof of Lemma 2.14]
Let a CNF formula \(\varphi \) have all variable indices below \(n\), and let \(a\) be an assignment satisfying \(\varphi \). Then some assignment agreeing with \(a\) on all indices below \(n\) satisfies the threaded clause transform of \(\varphi \) at cursor \(n\).
Given a flag (plain SAT or 3SAT mode) and a binary string \(z\), the verifier verdict: split \(|z|\) as \(i + (i+1)\) and reject outright if no such split exists; otherwise decode the first \(i\) bits of \(z\) to a formula \(\varphi \) and read an assignment off the remaining \(i + 1\) bits, and accept exactly when (in 3SAT mode) \(\varphi \) passes the width-three test and (in both modes) \(\varphi \) evaluates to true under that assignment.
Let \(x\) and \(u\) be binary words with \(|u| = |x| + 1\). Then the verifier verdict on the concatenation \(x \mathbin {+\! \! +} u\) equals the conjunction of the (mode-guarded) width-three test of the formula decoded from \(x\) with the evaluation of that formula under the assignment read off \(u\).
The single-bit function computing the plain-mode (SAT) verifier verdict on a binary string is polynomial-time computable: failed splits are rejected, failed parses take the accepting fallback branch, and all remaining inputs go through the safe evaluation pipeline.
The single-bit function computing the 3SAT-mode verifier verdict on a binary string is polynomial-time computable: the machine validates the entire syntax before running the width pass, and runs the evaluation pass only after width success; split failure rejects, while syntax failure accepts the empty fallback.
Let \(C\) be a clause, \(c \in \{ 0,1,2,3\} \) a counter value, and \(g\) a Boolean flag. Folding the width-pass transition over the serialization of \(C\), from the clause position with counter \(c\) and flag \(g\), returns the formula position with counter \(0\) and flag \(g \wedge [c + |C| \le 3]\), where \(|C|\) counts literal occurrences including repetitions.
Folding the width-pass transition over the serialization of a CNF formula \(\varphi \) from the initial state ends at the exact-end position with counter zero and the goodness flag conjoined with the Boolean width-three test of \(\varphi \); every clause is checked and the empty formula passes vacuously.
For every \(n \in \mathbb {N}\), every counter value \(c\), and every flag \(g\), folding the width-pass transition over the unary run \(1^n\) from the index position with counter \(c\) and flag \(g\) returns exactly the same state: index bits change neither the position, nor the counter, nor the flag.
Let \(\ell \) be a literal, \(c \in \{ 0,1,2,3\} \) a counter value, and \(g\) a Boolean flag. Folding the width-pass transition over the serialization of \(\ell \), from the clause position with counter \(c\) and flag \(g\), returns the clause position with counter \(\min (c + 1, 3)\) and flag \(g \wedge [c \lt 3]\), independently of the variable index and polarity of \(\ell \).
The transition of the separate width pass, which is run only after the complete syntax pass, acts on triples consisting of a position of the six-state syntax automaton, a saturated literal counter in \(\{ 0,1,2,3\} \), and a goodness flag. At formula level it follows the syntax automaton and resets the counter; at clause level a \(1\) (the start of a literal) moves to the index position, increments the saturated counter, and clears the flag if the counter was already \(3\), while a \(0\) returns to formula level with the counter reset; at every other position it follows the syntax automaton and leaves counter and flag unchanged. The pass never emits.
Given a configuration of a \((k+l)\)-tape machine whose state set is a product \(S \times D\), on an input word \(x\), the schedule projection is the configuration of a \(k\)-tape machine with state set \(S\) on the same input that keeps the \(S\)-component of the control state (halted projects to halted), the input-head position, and the first \(k\) work tapes with their head positions, and whose output stream is empty.
Let \(W\) be a Turing machine over the Boolean alphabet with \(k\) work tapes and let \(a \in \mathbb {N}\) be a padding parameter; the oblivious schedule machine built from \(W\) and \(a\) has \(k+3\) work tapes over the schedule alphabet. Given an input word \(x\) of schedule symbols, a base configuration \(c\) on \(x\), a controller phase \(q\) (possibly the halting state), an input-head position \(p \in \{ 0,\dots ,|x|+1\} \), integers \(b, u, g \in \mathbb {Z}\), and tape contents \(\beta , \mu , \gamma \colon \mathbb {Z}\to \Sigma \cup \{ \text{blank}\} \) (where \(\Sigma \) is the schedule alphabet), the associated setup configuration is the configuration with state \(q\) and input head at \(p\), whose first \(k\) work tapes and head positions agree with those of \(c\), whose last three (controller) tapes — the binary budget, unary macrostep counter, and sweep guide — have contents \(\beta , \mu , \gamma \) with heads at \(b, u, g\) respectively, and whose output agrees with that of \(c\).
Let \(n + 1 \le B\). From a setup configuration in the unary-start phase with unary head at cell \(B+1\) over the unary tape with \(B\) markers and guide head at cell \(0\) over the partial guide with \(n+1\) copied cells, exactly \(14B + 10\) steps produce the completed guide of radius \(3B\) and enter the macrostep-check phase with unary head at cell \(1\) and guide head at cell \(0\); the budget tape and head and the parked input head are unchanged.
From a setup configuration in append phase \(0\) with unary head at cell \(N+1\) over the unary tape with \(N\) markers, exactly \(a+2\) steps append exactly \(a+1\) unit markers and enter the borrow phase with incoming borrow set: the unary head is then at cell \(N+(a+1)+1\) over the unary tape with \(N+a+1\) markers, and all other tapes and heads are unchanged.
Let \(j \le a+1\), where \(a\) is the padding parameter. From a setup configuration in append phase \(0\) with unary head at cell \(N+1\) over the unary tape with \(N\) markers, exactly \(j\) steps reach append phase \(j\) with unary head at cell \(N+j+1\) over the unary tape with \(N+j\) markers; each step writes exactly one unit marker, and all other tapes and heads are unchanged.
Consider a setup configuration of the oblivious schedule machine built from a Boolean machine \(W\) with \(k\) work tapes and padding parameter \(a\), with input head at \(p\) and controller tapes \(\beta , \mu , \gamma \) with heads at \(b, u, g\). Applying a schedule action with next phase \(q'\), input-head move \(d \in \{ -1,0,+1\} \), and one (optional write, move) pair for each controller tape yields exactly the setup configuration with phase \(q'\), input head moved by \(d\) (clamped to \(\{ 0,\dots ,|x|+1\} \)), each controller tape updated by its optional write at the old head position, and each controller head shifted by its move; the first \(k\) work tapes and the output are untouched.
Let \(w\) be a bit word and \(c\) a Boolean incoming borrow, and recall the fixed-width borrow pass: it processes all of \(w\), subtracting the incoming borrow from its little-endian value while preserving its length, and returns an outgoing borrow flag together with the rewritten word. From a setup configuration in the borrow phase with flag \(c\) and budget head at cell \(1\) over the captured clock tape of \(w\), exactly \(|w|\) steps reach the borrow phase whose flag is the outgoing borrow, with budget head at \(1+|w|\) over the captured clock tape of the rewritten word; all other tapes and heads, and the parked input head, remain fixed.
From a setup configuration in the borrow phase with flag \(c\) and budget head at cell \(|w|+1\) (the blank cell after the captured clock tape of the bit word \(w\)), one step keeps every tape and head fixed and moves to the unary-start phase when \(c\) is true (the budget underflowed, i.e. is exhausted) and back to the budget-start phase otherwise.
Let \(w\) be a bit word of little-endian value \(v\) and let \(N \in \mathbb {N}\). From a setup configuration in the budget-start phase with budget head at \(|w|+1\) over the captured clock tape of \(w\) and unary head at \(N+1\) over the unary tape with \(N\) markers, exactly \((v+1)\, \bigl(2|w| + (a+1) + 4\bigr)\) steps reach the unary-start phase with unary head at \(N + (a+1)(v+1) + 1\) over the unary tape with \(N + (a+1)(v+1)\) markers, the budget tape being the captured clock tape of some bit word \(w'\) with \(|w'| = |w|\) (the word keeps its full width as its high bits become zero); the guide tape and head and the input head are unchanged.
For every \(j \in \mathbb {N}\): from a setup configuration in the budget-back phase with budget head at cell \(j\) over the captured clock tape of a bit word \(w\), exactly \(j+1\) steps reach the sentinel at the origin and move right once, entering the first append phase with budget head at cell \(1\); all tapes and all other heads are unchanged.
One complete round of the budget controller — full rewind, append block of \(a+1\) unary markers, fixed-width borrow pass, and end test — costs exactly \(2|w| + (a+1) + 4\) steps, independently of the bits of the budget word \(w\). Starting in the budget-start phase with budget head at \(|w|+1\) over the captured clock tape of \(w\) and unary head at \(N+1\) over the unary tape with \(N\) markers, the round ends with the captured clock tape of the borrowed word (incoming borrow set), the unary tape with \(N + a + 1\) markers, and phase unary-start or budget-start according to the borrow’s underflow flag.
From a setup configuration in the budget-start phase with budget head at cell \(|w|+1\) (one cell beyond the captured clock tape of the bit word \(w\)), exactly \(|w|+2\) steps reach the first append phase with budget head at cell \(1\); the cost depends only on the width of \(w\), not on its bits, and all tapes and all other heads are unchanged.
From a setup configuration in the copy-left phase with input head at position \(1\), guide head at cell \(0\), and entirely blank guide tape, exactly \(2|x| + 4\) steps — covering both virtual boundary cells and the return to the origin — reach the budget-start phase with input head parked at position \(|x|+1\), guide head at cell \(0\), and guide tape the partial guide with \(|x|+1\) copied cells; the budget and unary tapes and heads are unchanged, and the native input stays parked from then on.
From a setup configuration in the copy phase for \(|x|\) copied symbols, with input head parked at the right boundary \(|x|+1\) (reading blank) and guide head at \(|x|\), a single step enters the copy-return phase, leaves both heads in place, and closes the copied zone by writing one more mark, producing the partial guide with \(|x|+1\) cells.
From a setup configuration in the copy-left phase with guide head at \(0\) and entirely blank guide tape, exactly two steps of the oblivious schedule machine write the virtual left boundary (an interior mark at guide cell \(-1\)) and reach the first-copy phase with guide head back at \(0\) over the partial guide with \(0\) copied cells; the input head, the budget and unary tapes, and their heads are unchanged. This holds for every input-head position, including the empty input.
Let \(j \le |x|\). From a setup configuration in the first-copy phase with input head at position \(1\) and guide head at \(0\) over the partial guide with \(0\) copied cells, exactly \(j\) steps lead to the copy phase for \(j\) copied symbols with input head at \(j+1\), guide head at \(j\), and guide tape the partial guide with \(j\) copied cells; the budget and unary tapes and heads are unchanged. In particular the input and guide heads make exactly the same number of right moves, and no data value influences a movement.
For all \(n, j \in \mathbb {N}\): from a setup configuration in the copy-return phase with guide head at \(j\) over the partial guide with \(n+1\) copied cells, exactly \(j+1\) steps reach the budget-start phase with guide head at the origin (cell \(0\)); all tapes and all other heads are unchanged.
From a setup configuration in the macrostep-check phase with unary head at cell \(1\) over the unary tape with \(B\) markers and guide head at cell \(0\) over the completed guide of radius \(3B\), exactly \(B\, \bigl(6\, (3B) + 8\bigr) + 1\) steps execute one full sweep per unary marker and reach the halted configuration with unary head at cell \(B+1\) and guide head at cell \(0\); every tape, the budget head, and the input head are unchanged.
Let \(w\) be a bit word of little-endian value \(v\), and suppose the input satisfies \(|x| + 1 \le (a+1)(v+1)\), where \(a\) is the padding parameter. Consider a setup configuration in the reset-start phase with input head at an arbitrary position \(p\), budget head at cell \(|w|+1\) over the captured clock tape of \(w\), unary head at cell \(1\) over the unary tape with zero markers, and guide head at cell \(0\) over the blank guide tape. Then there is a bit word \(w'\) with \(|w'| = |w|\) such that after exactly
steps the machine is in the macrostep-check phase with input head parked at position \(|x|+1\), budget head at cell \(|w|+1\) over the captured clock tape of \(w'\), unary head at cell \(1\) over the unary tape with \((a+1)(v+1)\) markers, and guide head at cell \(0\) over the completed guide of radius \(3\, (a+1)(v+1)\): the binary counter retains its full width throughout.
In a setup configuration on input \(x\) with input head at position \(p \in \{ 0,\dots ,|x|+1\} \), the symbol read from the input tape is blank when \(p = 0\), and otherwise it is the \((p-1)\)-st symbol of \(x\) (in \(0\)-indexed order), which is blank when \(p - 1 \ge |x|\).
From layout phase \(0\) (of either direction) with unary head at cell \(B+1\) — the blank cell terminating the unary tape with \(B\) markers — and guide head at cell \(g\), one step writes an interior mark at guide cell \(g\), moves the guide head one further cell outward, keeps the unary head in place, and enters the right-edge or left-edge phase according to the direction.
Let \(\gamma \) be a guide tape carrying the origin mark exactly at cell \(0\), and suppose the unary tape reads a unit marker at the unary head. Then one step from layout phase \(i\) (of either direction) with guide head at \(\pm j\) over the \(j\)-cell filled version of \(\gamma \) leads to layout phase \(i+1 \bmod 3\) with guide head at \(\pm (j+1)\) over the \((j+1)\)-cell filled version, incrementing the unary head exactly when \(i = 2\).
Let \(B \gt 0\). From the first leftward layout phase with unary head at cell \(1\) over the unary tape with \(B\) markers and guide head at cell \(0\) over the right guide of radius \(3B\), exactly \(3B + 2\) steps complete the fixed guide and enter the unary-reset phase with unary head at \(B\), guide head at the left outer boundary \(-3B - 1\), and guide tape the completed guide of radius \(3B\).
Let \(\gamma \) be a guide tape carrying the origin mark exactly at cell \(0\), and let \(j \le 3B\). From layout phase \(0\) (of either direction) with unary head at cell \(1\) over the unary tape with \(B\) markers and guide head at cell \(0\) over \(\gamma \), exactly \(j\) steps lead to layout phase \(j \bmod 3\) with unary head at \(\lfloor j/3 \rfloor + 1\), guide head at \(\pm j\), and guide tape the \(j\)-cell filled version of \(\gamma \): the elapsed time equals the number of written guide cells, and the unary head moves once every three cells.
From layout-return phase \(0\) with unary head at cell \(B\) over the unary tape with \(B\) markers and guide head at cell \(3B\) (over an arbitrary guide tape), exactly \(3B+1\) steps reach the origin and enter the first leftward layout phase with unary head at cell \(1\) and guide head at cell \(0\); no tape is written.
Let \(j \le 3B\). From layout-return phase \(0\) with unary head at cell \(B\) over the unary tape with \(B\) markers and guide head at cell \(3B\) (over an arbitrary guide tape), exactly \(j\) steps lead to layout-return phase \(j \bmod 3\) with unary head at \(B - \lfloor j/3 \rfloor \) and guide head at \(3B - j\); no tape is written and all guide marks are unchanged.
Suppose a setup configuration is in layout-return phase \(i \in \{ 0,1,2\} \) and the unary tape does not read the origin sentinel at the unary head. Then one step moves to layout-return phase \(i+1 \bmod 3\), moves the guide head one cell left, decrements the unary head exactly when \(i = 2\), and writes nothing on any tape.
Let \(n + 1 \le B\). From the first rightward layout phase with unary head at cell \(1\) over the unary tape with \(B\) markers and guide head at cell \(0\) over the partial guide with \(n+1\) copied cells, exactly \(3B + 2\) steps reach radius \(3B\), install the right outer boundary marker, and enter the first layout-return phase with unary head at \(B\), guide head at \(3B\), and guide tape the right guide of radius \(3B\).
Consider a setup configuration in layout phase \(i \in \{ 0,1,2\} \) of either direction, and suppose the unary tape reads a unit marker at the unary head and the guide tape reads the origin mark at the guide head exactly when that head is at cell \(0\). Then one step moves to layout phase \(i+1 \bmod 3\), writes the guide mark at the guide head, moves the guide head one cell in the chosen direction, and advances the unary head by one exactly when \(i = 2\); everything else is unchanged. Thus the controller spends three guide cells per unary marker.
From a setup configuration in the reset-start phase with input head at an arbitrary position \(p \in \{ 0,\dots ,|x|+1\} \), running exactly \(\max (p,1)+1\) steps of the oblivious schedule machine leads to the setup configuration in the copy-left phase with input head at position \(1\); all controller tapes and heads are unchanged. (The reset begins with a mandatory left move, so even the right blank boundary and the empty input reach the left boundary before the final right move.)
Suppose the oblivious schedule machine is in a setup configuration in the reset-scan phase with input head at an interior position \(j \le |x|\); the controller tapes and heads are arbitrary. Then after exactly \(j\) steps the machine is in the same setup configuration except that the input head is at the left boundary (position \(0\)): the phase, all tapes, and all work heads are unchanged.
From a setup configuration in the right-edge phase with unary head at cell \(B+1\) over the unary tape with \(B\) markers and guide head at cell \(R+1\) over the partial guide with \(R+1\) copied cells, one step writes the right outer boundary marker at guide cell \(R+1\), moves the unary and guide heads one cell left (to \(B\) and \(R\)), and enters the first layout-return phase; the resulting guide tape is the right guide of radius \(R\).
For every \(j \in \mathbb {N}\): from a setup configuration in the start-center phase with guide head at cell \(-j\) over the completed guide of radius \(R\), exactly \(j+1\) steps reach the origin and enter the macrostep-check phase with guide head at cell \(0\); all tapes and all other heads are unchanged.
For every \(j \in \mathbb {N}\) and either rewind variant: from a setup configuration in the corresponding unary rewind phase with unary head at cell \(j\) over the unary tape with \(N\) markers, exactly \(j+1\) steps reach the successor phase — the start-center phase for the final rewind, the first rightward layout phase otherwise — with unary head at cell \(1\); no tape is written and all other heads are unchanged.
From a setup configuration in the unary-start phase with unary head at cell \(N+1\) over the unary tape with \(N\) markers, exactly \(N+2\) steps rewind the entire newly constructed unary segment and reach the first rightward layout phase with unary head at cell \(1\); all tapes and all other heads are unchanged.
Given a tape \(t \colon \mathbb {Z}\to A \cup \{ \text{blank}\} \), a cell \(z \in \mathbb {Z}\), and an optional write instruction, the resulting tape is \(t\) itself when no write is requested, and otherwise \(t\) updated at cell \(z\) with the requested (possibly blank) symbol.
Let \(M\) be a multi-tape Turing machine over the binary alphabet with a finite set \(Q\) of states and \(k\) work tapes, whose tape cells each hold a bit or the blank symbol \(\sqcup \). A snapshot of \(M\) is a triple \((q, a, b)\) consisting of a state \(q \in Q \cup \{ \bot \} \), where \(\bot \) marks a halted machine, an input symbol \(a \in \{ 0, 1, \sqcup \} \), and a tuple \(b = (b_0, \dots , b_{k-1})\) of work-tape symbols \(b_\tau \in \{ 0, 1, \sqcup \} \), one per work tape.
Let \(M\) be a binary multi-tape Turing machine with finite state set \(Q\) and \(k\) work tapes, let \(x \in \{ 0,1\} ^*\) and let \(t \in \mathbb {N}\). The snapshot of \(M\) on \(x\) at time \(t\) is the triple consisting of the state (\(\bot \) if halted), the symbol under the input head, and the \(k\) symbols under the work-tape heads, in the configuration reached by running \(M\) for \(t\) steps from its initial configuration on \(x\). Since a step leaves a halted configuration unchanged, this is defined at every time \(t\).
Let \(M\) be a binary multi-tape Turing machine that is oblivious, meaning that on any two inputs of the same length its input-head and work-head positions after \(t\) steps agree for every \(t\). Let \(x = x_1 \cdots x_n \in \{ 0,1\} ^n\) and \(t \in \mathbb {N}\), number the input cells \(0, 1, \dots , n+1\) with cell \(i \in \{ 1, \dots , n\} \) holding the \(i\)-th input bit, and let \(p\) be the position of the input head after \(t\) steps of \(M\) on the all-zeros input \(0^n\). Then the symbol under the input head after \(t\) steps of \(M\) on \(x\) is the blank \(\sqcup \) if \(p = 0\) or \(p \gt n\), and is the bit \(x_p\) otherwise.
Let \(M\) be a binary multi-tape Turing machine with finite state set \(Q\), \(k\) work tapes and transition function \(\delta \), let \(x \in \{ 0,1\} ^*\) and \(t \in \mathbb {N}\), and let \(q \in Q \cup \{ \bot \} \) (with \(\bot \) meaning halted), \(a\) and \(b = (b_0, \dots , b_{k-1})\) be the state, the symbol under the input head, and the symbols under the work heads after \(t\) steps of \(M\) on \(x\). Then the state of \(M\) on \(x\) after \(t + 1\) steps is \(\bot \) if \(q = \bot \), and otherwise is the successor state prescribed by the transition \(\delta (q, a, b)\).
Let \(M\) be a binary multi-tape Turing machine with finite state set \(Q\), transition function \(\delta \) and \(k\) work tapes indexed by \(\mathbb {Z}\), which is oblivious: on any two inputs of the same length its input-head and work-head positions after \(t\) steps agree for every \(t\). Let \(x \in \{ 0,1\} ^n\), \(t \in \mathbb {N}\) and \(\tau \in \{ 0, \dots , k-1\} \), and let \(w(s)\) be the position of the head of work tape \(\tau \) after \(s\) steps of \(M\) on \(0^n\). If no \(s \lt t\) has \(w(s) = w(t)\), then the symbol under head \(\tau \) after \(t\) steps of \(M\) on \(x\) is the blank \(\sqcup \). Otherwise let \(s\) be the largest such time and let \(q \in Q \cup \{ \bot \} \) (with \(\bot \) meaning halted), \(a\) and \(b = (b_0, \dots , b_{k-1})\) be the state, the input-head symbol and the work-head symbols after \(s\) steps of \(M\) on \(x\); then the symbol under head \(\tau \) after \(t\) steps of \(M\) on \(x\) is \(b_\tau \) if \(q = \bot \), and otherwise is the symbol that \(\delta (q, a, b)\) writes on tape \(\tau \), or \(b_\tau \) if that transition writes nothing on tape \(\tau \).
Let \(M\) be a binary multi-tape Turing machine with initial state \(q_0\) and \(k\) work tapes, and let \(x \in \{ 0,1\} ^*\). Before any step is taken, \(M\) on input \(x\) is in state \(q_0\), the symbol under its input head is the first bit \(x_1\) of \(x\) (or the blank \(\sqcup \) if \(x\) is empty), and every one of its \(k\) work heads reads the blank \(\sqcup \).
Let \(M\) be a binary machine that decides a language \(L\) over the binary alphabet within time \(a \cdot T(n)\) on inputs of length \(n\). Then for every input \(x\), after running \(M\) for \((a + 1)(T(|x|) + 1)\) steps from its initial configuration on \(x\), the last bit of the output word, defaulting to false, equals the indicator of the membership of \(x\) in \(L\).
Let a configuration of a binary machine with \(k\) work tapes on an input \(x\) be given, and let \(z\) be an integer offset. The head-centered payload vector at \(z\) assigns to each work tape the symbol of that tape at signed distance \(z\) from its head, with ordinary tag \(0\), and assigns to the final coordinate the virtual input payload of \(x\) at the input head position minus one plus \(z\); thus the last virtual tape is the copied input and the earlier ones are the source work tapes.
Let \(c\) be a configuration of a binary machine with \(k\) work tapes, let \(\alpha \) be an action, let \(z\) be an integer offset, and let \(i\) be one of the \(k+1\) virtual tape coordinates. Then the head-centered payload of the configuration obtained by applying \(\alpha \) to \(c\), at offset \(z\) and coordinate \(i\), equals the written payload of \(c\) and \(\alpha \) at offset \(z\) plus the shift direction of coordinate \(i\): writing at the old origin and then shifting each virtual tape by its displacement produces exactly the payloads of the successor configuration, including stationary moves, boundary attempts, and blank writes.
Let \(M\) be a binary machine, \(x\) an input, and \(t, B\) natural numbers with \(t \le B\) and \(|x| + 1 \le B\). Then for every integer coordinate \(z\) with \(z \lt -3B\) or \(z \gt 3B\) and every virtual tape coordinate, the head-centered payload of the configuration reached by running \(M\) for \(t\) steps from its initial configuration on \(x\) is the blank payload. This bound uses the initialized run, not arbitrary starting configurations.
For a configuration of a binary machine with \(k\) work tapes, the read vector at a macrostep origin assigns to each work tape the symbol under its head with ordinary tag \(0\), and to the final coordinate the symbol read by the input head paired with the boundary tag of the configuration.
Given a configuration of a machine with \(k\) work tapes and an action, the shift vector assigns to each work tape the direction prescribed by the action for that tape, and to the final coordinate the clipped direction of the action’s input movement at the current input head position.
For a binary machine \(M\) and a configuration \(c\) of \(M\), the totalized action is the identity action (no writes, no output, no head movement, halted successor) when \(c\) is halted, and the transition action of \(M\) on the current input and work-tape reads of \(c\) otherwise; halting is thereby interpreted as an idle macrostep.
For every binary machine \(M\) and every configuration \(c\) of \(M\), applying the totalized action of \(M\) at \(c\) to \(c\) yields exactly one machine step of \(M\) from \(c\), including the idle step performed after \(M\) has halted.
Given a configuration of a machine with \(k\) work tapes, an action, and an integer offset \(z\), the written payload vector describes the tapes after performing the action’s writes at the old origin but before any shifting: at offset \(0\) each work-tape entry is the symbol written by the action (or the symbol under the head when the action writes nothing), at every other offset it is the untouched tape content, and the input entries are read-only, always given by the virtual input payload.
For a bit string \(v\), let \(\mathrm{split}(v)\) be the prefix of \(v\) preceding its last occurrence of \(1\), undefined when \(v\) contains no \(1\). Then for every \(n \in \mathbb {N}\), \(\mathrm{split}(1^{n+1})\) is defined and equals \(1^n\), where \(1^k\) is the word of \(k\) ones.
Let \(M\) be a binary multi-tape Turing machine with finite state set \(Q\), \(k\) work tapes and transition function \(\delta \), and let \((q, a, b)\) be a snapshot of \(M\), consisting of a state \(q \in Q \cup \{ \bot \} \) (\(\bot \) meaning halted), an input symbol \(a \in \{ 0, 1, \sqcup \} \) and work-tape symbols \(b = (b_0, \dots , b_{k-1})\) in \(\{ 0, 1, \sqcup \} \). Its successor state is \(\bot \) if \(q = \bot \), and otherwise the successor state in \(Q \cup \{ \bot \} \) prescribed by the transition \(\delta (q, a, b)\).
The partial map \(\mathrm{strip}\) on binary strings that deletes the last \(1\) together with the run of \(0\)s following it. It is defined recursively: \(\mathrm{strip}(\varepsilon )\) is undefined, and for a bit \(b\) and a string \(v\), \(\mathrm{strip}(bv) = b\, \mathrm{strip}(v)\) when \(\mathrm{strip}(v)\) is defined, while if \(\mathrm{strip}(v)\) is undefined then \(\mathrm{strip}(1v) = \varepsilon \) and \(\mathrm{strip}(0v)\) is undefined. Thus \(\mathrm{strip}(v)\) is the prefix of \(v\) preceding its last \(1\), and is undefined when \(v\) contains no \(1\).
Let \(\mathrm{strip}\) be the partial map on binary strings defined by: \(\mathrm{strip}(\varepsilon )\) is undefined, and for a bit \(b\) and a string \(v\), \(\mathrm{strip}(bv) = b\, \mathrm{strip}(v)\) when \(\mathrm{strip}(v)\) is defined, while otherwise \(\mathrm{strip}(1v) = \varepsilon \) and \(\mathrm{strip}(0v)\) is undefined. Then for every \(k \in \mathbb {N}\), \(\mathrm{strip}(0^k)\) is undefined.
Let \(\mathrm{strip}\) be the partial map on binary strings defined by: \(\mathrm{strip}(\varepsilon )\) is undefined, and for a bit \(b\) and a string \(v\), \(\mathrm{strip}(bv) = b\, \mathrm{strip}(v)\) when \(\mathrm{strip}(v)\) is defined, while otherwise \(\mathrm{strip}(1v) = \varepsilon \) and \(\mathrm{strip}(0v)\) is undefined. Then for every binary string \(u\) (possibly empty, possibly containing \(1\)s) and every \(k \in \mathbb {N}\), \(\mathrm{strip}(u\, 1\, 0^k) = u\).
Let \(\mathrm{strip}\) be the partial map on binary strings defined by: \(\mathrm{strip}(\varepsilon )\) is undefined, and for a bit \(b\) and a string \(v\), \(\mathrm{strip}(bv) = b\, \mathrm{strip}(v)\) when \(\mathrm{strip}(v)\) is defined, while otherwise \(\mathrm{strip}(1v) = \varepsilon \) and \(\mathrm{strip}(0v)\) is undefined. Then for every binary string \(v\): \(\mathrm{strip}(v)\) is undefined if and only if \(v = 0^{|v|}\) consists only of zeros; and for every binary string \(u\), \(\mathrm{strip}(v) = u\) if and only if \(v = u\, 1\, 0^k\) for some \(k \in \mathbb {N}\).
Present a tape as a zipper with frontier \(z \in \mathbb {Z}\), a list \(l\) of the (possibly blank) cells to the left of \(z\) listed nearest-first, and a list of the cells at and to the right of \(z\). If the right list decomposes as \(w \mathbin {+\! \! +} r\), then the same tape is presented by the zipper with frontier \(z + |w|\), left list \(\mathrm{reverse}(w) \mathbin {+\! \! +} l\), and right list \(r\).
For every binary string \(x\), \(x\) lies outside \(\mathrm{TAUTOLOGY}\) if and only if there exists a binary word \(u\) of length exactly \(|x| + 1\) such that the DNF decoded from \(x\) evaluates to false under the assignment \(\alpha _u\) that reads variable \(v\) off the \(v\)-th bit of \(u\) (false for \(v \ge |u|\)).
Let \(C\) be a list of literals whose variable indices are all below \(|u|\) for a word \(u\), let \(\sigma \) be its serialization, and let the input be \(x = p \mathbin {+\! \! +} \mathrm{pair}(\sigma \mathbin {+\! \! +} r, u)\). Started in the clause-level evaluation state with accumulator \(c\), \(|p|\) bits consumed, work word \(u\), and head \(0\), the non-tautology verifier within some \(t \le 3|\sigma |\) steps has consumed \(|p| + 2|\sigma |\) bits, has head \(0\), and is in the verdict state with bit \(0\) if \(c\) holds and every literal of \(C\) is satisfied by the assignment read off \(u\), and in the formula-level evaluation state otherwise.
Let \(M\) and \(U\) be finite multi-tape Turing machines over the binary alphabet, let \(f, g\) be functions on binary strings, and let \(T_1, T_2 : \mathbb {N} \to \mathbb {N}\). If \(M\) computes \(f\) within \(T_1\), and for every input \(x\) the machine \(U\) run on \(f(x)\) halts within \(T_2(|x|)\) steps with output \(g(x)\), then some finite machine computes \(g\) within time \(n \mapsto 2\, T_1(n) + T_2(n) + 2\).
Let the input be \(x = p \mathbin {+\! \! +} u\). Started in the certificate-copy state with \(|p|\) input bits consumed, empty work tape, work head \(0\), and empty output, for every \(j \le |u|\) the non-tautology verifier after exactly \(j\) steps is still in the copy state with \(|p| + j\) bits consumed, the first \(j\) bits of \(u\) on the work tape, the work head at cell \(j\), and empty output.
Let \(\psi \) be a DNF formula and let \(a, b : \mathbb {N} \to \{ 0,1\} \) be assignments with \(a(v) = b(v)\) for every \(v\) below the variable bound of \(\psi \) (one more than its largest variable index). Then \(\psi \) evaluates to the same truth value under \(a\) and under \(b\).
Let \(x\) be an input with \(i + 2 \le |x|\) and \(i\)-th bit \(b\), let \(u\) be a word, \(h\) an integer, and \(q\) an evaluation state. Suppose the evaluation action for \(q\) on bit \(b\), with the work cell at \(h\) of the tape holding \(u\), is the silent move advancing the input head, moving the work head by \(e\), and entering \(q'\). Then from the non-tautology verifier’s evaluation configuration carrying \(q\) with \(i\) bits consumed, work word \(u\), and head \(h\), two steps reach state \(q'\) with \(i+2\) bits consumed and head \(h + e\).
Let \(x\) be a binary string on which the exact-consumption formula parser succeeds, and let \(u\) be any binary word. Then for some \(t \le 4(|\mathrm{pair}(x,u)| + 2)\), running the non-tautology verifier for \(t\) steps from its initial configuration on \(\mathrm{pair}(x,u)\) reaches the formula-level evaluation state with no input bits consumed, \(u\) on the work tape from cell \(0\), the work head at cell \(0\), and empty output.
Let \(\varphi \) be a formula whose variable indices are all below \(|u|\) for a word \(u\), let \(\sigma \) be its serialization, and let the input be \(x = p \mathbin {+\! \! +} \mathrm{pair}(\sigma , u)\). Started in the formula-level evaluation state with \(|p|\) bits consumed, work word \(u\), and head \(0\), the non-tautology verifier halts within some \(t \le 3|\sigma | + 1\) steps, with work head \(0\) and output the single bit that negates the truth value of \(\varphi \) read as a DNF under the assignment \(\alpha _u\) reading variable \(v\) off the \(v\)-th bit of \(u\) (false beyond \(|u|\)).
For every input \(x\), the initial configuration of the non-tautology verifier on \(x\) is the configuration in the syntax state carrying the formula-level scanner state, with no input bits consumed, an empty work tape, work head at cell \(0\), and empty output.
Let \(\varphi \) be a CNF formula, \(C\) a clause of \(\varphi \), and \(\ell = (v, b)\) a literal of \(C\). Then \(v\) is strictly less than the variable bound of \(\varphi \), i.e. one more than the largest variable index occurring in \(\varphi \).
Let \(\ell = (v, b)\) be a literal and \(u\) a word with \(v \lt |u|\), and let the input be \(x = p \mathbin {+\! \! +} \mathrm{pair}(\sigma \mathbin {+\! \! +} r, u)\) where \(\sigma \) is the serialization of \(\ell \). Started in the clause-level evaluation state with accumulator \(c\), \(|p|\) bits consumed, work word \(u\), and head \(0\), the non-tautology verifier after exactly \(3v + 7\) steps is in the clause-level state with accumulator \(c \wedge (\alpha _u(v) = b)\), where \(\alpha _u\) reads \(v\) off the \(v\)-th bit of \(u\), with \(|p| + 2|\sigma |\) bits consumed and head \(0\).
Let \(x\) be a binary string on which the exact-consumption formula parser fails, and let \(u\) be any binary word. Then the non-tautology verifier, run on \(\mathrm{pair}(x, u)\), halts within \(2|x| + 3\) steps with output the single bit \(1\).
Let \(x\) and \(u\) be binary words with \(|u| = |x| + 1\). The non-tautology verifier, run on \(\mathrm{pair}(x, u)\), halts within \(10(|\mathrm{pair}(x,u)| + 1)\) steps with output the single bit that negates the truth value of the DNF decoded from \(x\) under the assignment \(\alpha _u\) reading variable \(v\) off the \(v\)-th bit of \(u\) (false beyond \(|u|\)).
For every binary string \(z\), let \(z'\) be the pair encoding of the first \(i\) bits of \(z\) with the rest if \(|z| = 2i + 1\), and the empty word if \(|z|\) is even. The non-tautology verifier, run on \(z'\), halts within \(40(|z| + 1)\) steps with output the single bit that is \(0\) if \(|z|\) is even and otherwise negates the truth value of the DNF decoded from the first \(i\) bits of \(z\) under the assignment read off the remaining bits.
Let \(V\) be the set of binary strings \(z\) with \(|z| = i + (i + 1)\) for some \(i\) such that the DNF decoded from the first \(i\) bits of \(z\) is false under the assignment reading variable \(v\) off the \(v\)-th of the remaining bits (false beyond them). If \(V\) belongs to \(\mathrm{P}\), then \(\mathrm{TAUTOLOGY}\) belongs to coNP.
Let \(x\) be an input with \(i + 1 \le |x|\). Applying the silent action that moves the input head right, moves the work head by \(e\), and enters state \(q'\) to a non-tautology verifier configuration on \(x\) in any state with \(i\) input bits consumed, work word \(u\), work head \(h\), and output \(o\) yields the configuration in state \(q'\) with \(i+1\) bits consumed, the same \(u\) and \(o\), and work head \(h + e\).
Let \(x\) be an input and \(i \le |x|\). Applying the silent action that keeps the input head fixed, moves the work head by \(e\), and enters state \(q'\) to a non-tautology verifier configuration on \(x\) in any state with \(i\) input bits consumed, work word \(u\), work head \(h\), and output \(o\) yields the configuration in state \(q'\) with the same index, \(u\), and \(o\), and work head \(h + e\).
For every DNF formula \(\psi \), the CNF obtained from \(\psi \) by flipping the polarity of every literal has the same variable bound as the term list of \(\psi \) viewed as a CNF, where the variable bound of a formula is one more than its largest variable index (and \(0\) if no literal occurs).
For all binary words \(a\), \(b\), \(u\), the pair encoding satisfies \(\mathrm{pair}(a \mathbin {+\! \! +} b, u) = \overline{a} \mathbin {+\! \! +} \mathrm{pair}(b, u)\), where \(\overline{a}\) is the word obtained from \(a\) by doubling every bit.
Let \(x\) be an input and \(i \le |x|\). In every configuration of the non-tautology verifier on \(x\) whose input head has consumed \(i\) bits, the symbol under the input head is the \(i\)-th bit of \(x\) when \(i \lt |x|\) and the blank when \(i = |x|\), regardless of state, work tape, and output.
Let \(f\) be a fuel value and \(x\) a binary string with \(|x| \le f\). The six-state formula syntax scan from the clause-level state accepts \(x\) if and only if the clause parser with fuel \(f\) succeeds on \(x\) with some remainder \(r\) and the scan from the formula-level state accepts \(r\); when the clause parser fails, the scan rejects.
For every state \(q\) of the six-state formula syntax scanner with transition \(\delta \), every bit \(b\), and every binary string \(x\), the acceptance bit of the scan from \(q\) on \(b :: x\) equals the acceptance bit of the scan from \(\delta (q, b)\) on \(x\).
Let \(x\) be an input with \(i + 2 \le |x|\) whose bits at positions \(i\) and \(i + 1\) both equal \(b\), and let \(s\) be a formula-scanner state with transition \(\delta \). Started in the syntax state carrying \(s\) with \(i\) input bits consumed, empty work tape, work head \(0\), and empty output, the non-tautology verifier reaches, after exactly two steps, the syntax state carrying \(\delta (s, b)\) with \(i + 2\) bits consumed and all else unchanged.
Let \(f\) be a fuel value and \(x\) a binary string with \(|x| \le f\). The six-state formula syntax scan from the formula-level state accepts \(x\) if and only if the formula parser with fuel \(f\) succeeds on \(x\) with an empty remainder.
For every binary string \(x\), the string \(1 :: x\) is accepted by the six-state formula syntax scan from the clause-level state if and only if the literal parser succeeds on \(1 :: x\) with some remainder \(r\) and \(r\) is accepted by the scan from the clause-level state; when the literal parser fails, the scan rejects.
Let \(a, u, p\) be binary words, \(x = p \mathbin {+\! \! +} \mathrm{pair}(a, u)\), and \(s\) a formula-scanner state. Started in the syntax state carrying \(s\) with \(|p|\) input bits consumed, empty work tape, head \(0\), and empty output, the non-tautology verifier after exactly \(2|a| + 2\) steps has consumed \(|p| + 2|a| + 2\) bits, with work tape and output still empty, and is in the certificate-copy state if folding the scanner transition over \(a\) from \(s\) ends in the exact-end state, and in the verdict state with bit \(1\) otherwise.
Let \(x\) be an input with \(i + 2 \le |x|\) whose bits at positions \(i\) and \(i+1\) are \(0\) and \(1\), and let \(s\) be a formula-scanner state. Started in the syntax state carrying \(s\) with \(i\) input bits consumed, empty work tape, head \(0\), and empty output, after exactly two steps the non-tautology verifier has consumed \(i+2\) bits and is in the certificate-copy state if \(s\) is the exact-end state, and in the verdict state with bit \(1\) otherwise.
For every binary string \(x\), write \(x = 1^k s\) where \(s\) is the suffix left by the leading-ones scanner. The six-state formula syntax scan from the unary-index state accepts \(x\) if and only if \(s = 0\, b\, r\) for some bit \(b\) and string \(r\) and the scan from the clause-level state accepts \(r\); if \(s\) has fewer than two bits or starts with \(1\), the scan rejects.
Let \(k \in \mathbb {N}\), let \(r, u, p\) be binary words, and let the input be \(x = p \mathbin {+\! \! +} \mathrm{pair}(1^k \mathbin {+\! \! +} r, u)\). Started in the unary-index evaluation state with accumulator \(c\), \(|p|\) input bits consumed, work word \(u\), and work head \(h\), the non-tautology verifier after exactly \(2k\) steps is in the same state with \(|p| + 2k\) bits consumed and work head \(h + k\).
Let \(x\) be an input, \(i \le |x|\), \(u\) a word, \(h\) an integer, and \(b\) a bit. From the non-tautology verifier’s configuration in the verdict state with bit \(b\), \(i\) input bits consumed, work word \(u\), work head \(h\), and empty output, one step halts with output exactly the single bit \(b\), all other components unchanged.
Let \(x\) and \(u\) be binary words with \(|u| = |x| + 1\), and let \(V\) be the set of binary strings \(z\) with \(|z| = i + (i+1)\) whose first \(i\) bits decode to a DNF that is false under the assignment read off the remaining bits. Then \(x \mathbin {+\! \! +} u \in V\) if and only if the DNF decoded from \(x\) evaluates to false under the assignment \(\alpha _u\) reading variable \(v\) off the \(v\)-th bit of \(u\) (false beyond \(|u|\)).
The set of binary strings \(z\) such that \(|z| = i + (i + 1)\) for some \(i\) and the DNF decoded from the first \(i\) bits of \(z\) evaluates to false under the assignment reading variable \(v\) off the \(v\)-th of the remaining \(i + 1\) bits (false beyond them) belongs to \(\mathrm{P}\).
Let \(q\) and \(d\) be control states of the non-tautology verifier such that, in state \(q\), it silently moves the work head left with the input head fixed whenever the work cell is nonblank, and otherwise moves the work head right and enters \(d\). Then for every input \(x\), index \(i \le |x|\), work word \(u\), and \(n \le |u|\), starting in state \(q\) with work head at cell \(n - 1\), exactly \(n + 1\) steps reach state \(d\) with work head at cell \(0\), all other components unchanged.
Given an input \(x\), an optional control state \(q\) (absent meaning halted), an index \(i \le |x|\), a word \(u\), an integer \(h\), and an output word (empty by default), the indexed configuration is the configuration of the one-work-tape non-tautology verifier on \(x\) with state \(q\), the input head having consumed \(i\) input bits, the work tape holding \(u\) from cell \(0\) (blank elsewhere), the work head at cell \(h\), and the given output.
The control states of the non-tautology verifier machine are: syntax states carrying a formula-scanner state, either before the first copy of a doubled input bit or holding the first copy already read; certificate-copy and copy-back states; two input-rewind states; evaluation states carrying an evaluation state, either before a doubled bit or holding its first copy; a certificate-rewind state carrying the term accumulator; and a verdict state carrying the output bit.
For every input \(x\), applying the zero-work-tape action that moves the input head by \(d\), enters \(q'\), and optionally emits \(o\) to the configuration with state \(q\), head position \(p\), and output \(w\) yields the configuration with state \(q'\), head position \(p\) moved by \(d\) (clamped at the boundary blanks), and output \(w\) followed by \(o\).
For every clause \(C\) and binary string \(r\), the polarity-flipping scan from the clause-level state of the formula syntax scanner maps the serialization of \(C\) followed by \(r\) to the serialization of the clause obtained from \(C\) by negating every literal’s polarity, followed by the polarity-flipping scan of \(r\) from the formula-level state.
Let the input be \(x = p \mathbin {+\! \! +} r\), let \(q\) be a formula-scanner state, and let \(o\) be an output word. Started in the copying state carrying \(q\) at head position \(|p| + 1\) with output \(o\), the dual transducer reaches, in exactly \(|r|\) steps and without visiting the anchor state before then, the copying state carrying the fold of the scanner transition over \(r\) from \(q\), at position \(|x| + 1\), with output \(o\) followed by the polarity-flipping scan of \(r\) from \(q\).
For every input \(x\), output word \(o\), and formula-scanner state \(q\), the configuration reached by one step of the dual transducer from the copying state carrying \(q\) at the right boundary position \(|x| + 1\) with output \(o\) is the rewind state with flag true at position \(|x|\) with output still \(o\), and the anchor state is not visited.
Let \(x\) be a binary string on which the six-state formula syntax scan from formula level does not end in the exact-end state. Started in the formula-level scanning state at head position \(1\) with empty output, the dual transducer reaches the anchor state at position \(1\) with output the single bit \(0\), in exactly \(2|x| + 2\) steps and without visiting the anchor before then.
For every literal \((v, b)\) and binary string \(r\), the polarity-flipping scan from the clause-level state of the formula syntax scanner maps the serialization of \((v, b)\) followed by \(r\) to the serialization of \((v, \neg b)\) followed by the polarity-flipping scan of \(r\) from the clause-level state.
For every binary string \(x\), the guarded dual output of \(x\) — the polarity-flipping scan of \(x\) if \(x\) passes the formula syntax scan, and the word \(0\) otherwise — equals the serialization of the De Morgan dual (every literal negated, read as a DNF) of the CNF formula decoded from \(x\).
Let \(x\) be an input and \(i \le |x|\). In every configuration of the dual transducer on \(x\) with input head at position \(i + 1\), the symbol read is the \(i\)-th bit of \(x\) when \(i \lt |x|\) and the right boundary blank when \(i = |x|\).
Let \(x\) be an input, \(o\) an output word, \(f\) a flag, and \(p \le |x|\). Started in the rewind state with flag \(f\) at head position \(p\) and output \(o\), the dual transducer reaches head position \(1\) with output \(o\), in the anchor state if \(f\) is true and in the formula-level copying state otherwise, in exactly \(p + 1\) steps and without visiting the anchor before then.
For every binary string \(x\), let \(c\) be the dual transducer’s configuration on \(x\) in the anchor state with input head at position \(1\) and empty output. There is a time \(t\) with \(0 \lt t \le 4|x| + 5\) such that the configurations at all times \(j\) with \(0 \lt j \lt t\) are not in the anchor state, and the configuration at time \(t\) equals \(c\) except that its output is the guarded dual output of \(x\): the polarity-flipping scan of \(x\) if \(x\) passes the formula syntax scan, and the single bit \(0\) otherwise.
Let the input be \(x = p \mathbin {+\! \! +} r\), let \(q\) be a formula-scanner state, and let \(o\) be an output word. Started in the scanning state carrying \(q\) at head position \(|p| + 1\) with output \(o\), the dual transducer reaches the scanning state carrying the fold of the scanner transition over \(r\) from \(q\), at position \(|x| + 1\) with output still \(o\), in exactly \(|r|\) steps and without visiting the anchor state before then.
Let \(c, d, e\) be configurations of the dual transducer. If running \(a\) steps from \(c\) reaches \(d\) without visiting the anchor state before time \(a\), and running \(b\) steps from \(d\) reaches \(e\) without visiting the anchor before time \(b\), then running \(a + b\) steps from \(c\) reaches \(e\) without visiting the anchor before time \(a + b\).
Let \(c\) and \(d\) be configurations of the dual transducer such that one step takes \(c\) to \(d\) and \(c\) is not in the anchor state. Then one step from \(c\) reaches \(d\) with no anchor visit before it, i.e. \(c\) to \(d\) is an anchor-free segment of length \(1\).
For every CNF formula \(\varphi \), the polarity-flipping scan from the formula-level state of the formula syntax scanner maps the serialization of \(\varphi \) to the serialization of its De Morgan dual, the DNF obtained from \(\varphi \) by negating every literal in place.
For every \(n \in \mathbb {N}\) and binary string \(r\), the polarity-flipping scan from the unary-index state of the formula syntax scanner maps \(1^n \mathbin {+\! \! +} r\) to \(1^n\) followed by the polarity-flipping scan of \(r\) from the unary-index state.
Let \(x\) be a binary string on which the six-state formula syntax scan from formula level ends in the exact-end state. Started in the formula-level scanning state at head position \(1\) with empty output, the dual transducer reaches the anchor state at position \(1\) with output the polarity-flipping scan of \(x\), in exactly \(4|x| + 4\) steps and without visiting the anchor before then.
Given an input-head displacement \(d \in \{ -1, 0, 1\} \), a control state \(q\), and an optional output bit \(o\) (absent by default), the transducer action is the action of a machine with no work tapes that moves the input head by \(d\), emits \(o\) if present, and enters \(q\).
Given an input \(x\), a control state \(q\), an input-head position \(p \in \{ 0, \dots , |x| + 1\} \) (positions \(0\) and \(|x|+1\) being the boundary blanks), and an output word \(o\), the transducer configuration is the configuration of the zero-work-tape dual transducer on \(x\) with state \(q\), input head at \(p\), and output \(o\).
The control states of the dual transducer are an anchor state, scanning states and copying states each carrying a formula-scanner state, and rewind states carrying a flag that selects whether the rewind ends at the anchor or starts the copying pass.
Given a binary string \(x\), the guarded dual output is the polarity-flipping scan of \(x\) from the formula-level state when the six-state formula syntax scan of \(x\) from formula level ends in the exact-end state, and the one-bit word \(0\) (the serialization of the empty formula) otherwise.
Given a state \(q\) of the six-state formula syntax scanner and a binary string \(x\), the polarity-flipping scan is the string obtained by running the scanner over \(x\) from \(q\) and outputting each bit as read, except that a bit read in the polarity state is flipped.
Given configurations \(c\) and \(d\) of the dual transducer on an input \(x\) and a time \(t \in \mathbb {N}\), the segment predicate holds when running the transducer for \(t\) steps from \(c\) reaches exactly \(d\), and for every \(j \lt t\) the configuration after \(j\) steps from \(c\) is not in the anchor state.
The dual transducer is a finite Turing machine over the binary alphabet with no work tapes. From its anchor state it scans the whole input silently with the six-state formula syntax scanner. If the scan ends in the exact-end state, it rewinds and makes a second pass emitting every input bit, flipped exactly when read in the polarity state; otherwise it emits the single bit \(0\). In both cases it then rewinds to the first input position and re-enters the anchor state.
The evaluation states of the non-tautology verifier are formula level, and the states clause level, unary index, and polarity, which each carry a Boolean accumulator recording whether every literal of the current term read so far is satisfied.
Given an evaluation state, a formula bit \(b\), and the work cell \(w\) under the certificate head (blank read as \(0\)), the evaluation action is a silent move that advances the input head and acts as follows. At formula level, \(1\) opens a term with accumulator true and \(0\) enters the accepting verdict. At clause level with accumulator \(c\), \(1\) starts a literal and \(0\) ends the term, entering the rejecting verdict if \(c\) is true and returning to formula level otherwise. In the unary index, each \(1\) advances the certificate head and \(0\) moves to polarity; the polarity bit \(b\) sets the accumulator to \(c \wedge (w = b)\), moves the certificate head left, and enters the certificate-rewind state.
Given an input-head displacement \(d\) and a work-head displacement \(e\) in \(\{ -1, 0, 1\} \) and a control state \(q\), the silent move is the one-work-tape machine action that moves the input head by \(d\) and the work head by \(e\), writes nothing, emits nothing, and enters \(q\).
The language \(\mathrm{TAUTOLOGY}\) is the set of binary strings \(x\) whose decoded formula, read as a DNF (an OR of ANDs of literals), evaluates to true under every assignment \(a : \mathbb {N} \to \{ 0,1\} \) [AB09, Example 2.21, DNF fragment]. Since decoding is total with the empty formula as fallback, and the empty disjunction is false, no non-well-formed string belongs to \(\mathrm{TAUTOLOGY}\).
The transition function \(\delta \) of the six-state syntax scanner for the formula serialization grammar is the following: at formula level, \(1\) goes to clause level and \(0\) to exact end; at clause level, \(1\) goes to unary index and \(0\) back to formula level; in the unary index, \(1\) stays and \(0\) goes to polarity; from polarity, any bit returns to clause level; from exact end or error, any bit goes to error.
The non-tautology verifier is a finite Turing machine over the binary alphabet with one work tape. On an input \(\mathrm{pair}(x, u)\) it reads the doubled bits of \(x\) two at a time through the six-state formula syntax scanner; at the separator it emits \(1\) and halts if \(x\) is not well formed, and otherwise copies \(u\) onto the work tape, rewinds both heads, and evaluates the DNF serialized by \(x\) term by term against the certificate assignment on the work tape, finally emitting the negation of its truth value. An input ending before a separator, or a doubled pair \(10\), makes it emit \(0\); no bit is emitted before the final verdict.
The verifier language is the set of binary strings \(z\) such that \(|z| = i + (i + 1)\) for some \(i\) and the DNF decoded from the first \(i\) bits of \(z\) evaluates to false under the assignment reading variable \(v\) off the \(v\)-th bit of the remaining \(i + 1\) bits (false beyond them).
Given a binary string \(z\), the verifier bit is false when \(|z|\) cannot be written as \(i + (i + 1)\); otherwise, for the unique such \(i\), it is the negation of the truth value of the DNF decoded from the first \(i\) bits of \(z\) under the assignment that reads variable \(v\) off the \(v\)-th bit of the remaining suffix of \(z\) (false beyond its end).
Let \(L \subseteq \{ 0,1\} ^*\) be a language in \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time. Then the map sending a bit string \(z\) to the one-bit word \(\mathbf{1}_L(z)\), which is \(1\) if \(z \in L\) and \(0\) otherwise, is polynomial-time computable, i.e. some finite binary Turing machine computes it within \(C(n+1)^c\) steps on every input of length \(n\), for some constants \(C, c \in \mathbb {N}\).
A function \(T : \mathbb {N} \to \mathbb {N}\) is time constructible if \(n \le T(n)\) for every \(n\), and there exist a positive integer constant \(c\) and a finite Turing machine over the binary alphabet that, on every input word \(x \in \{ 0,1\} ^{*}\), computes the binary representation of \(T(|x|)\) (least-significant bit first, with no redundant leading zeros, so \(0\) is represented by the empty word) within at most \(c \cdot (T(|x|) + 1)\) steps.
The identity function \(n \mapsto n\) on the natural numbers is time constructible: trivially \(n \le n\) for all \(n\), and there are a positive integer constant \(c\) and a finite Turing machine over the binary alphabet that, on every input word \(x \in \{ 0,1\} ^{*}\), outputs the binary representation of \(|x|\) (least-significant bit first, empty for \(|x| = 0\)) within at most \(c \cdot (|x| + 1)\) steps.
For all \(C, c \in \mathbb {N}\) with \(C \gt 0\), the function \(T(n) = C(n+1)^{c+1}\) is time constructible: \(n \le T(n)\) for every \(n\), and for some constant \(a \gt 0\) a finite binary Turing machine outputs, on every input \(x\), the little-endian binary representation of \(T(|x|)\) within \(a\, (T(|x|) + 1)\) steps.
Let \(c\) be an effective machine-representation scheme: it assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \), and comes with a canonizer machine that, on every input \(\alpha \), outputs the canonical serialization of \(M_\alpha \) within \(T_c(|\alpha |)\) steps. Then there is a single finite binary Turing machine \(U\) such that for all bit strings \(\alpha , x\) and every \(t \in \mathbb {N}\), writing \(S = (3|\alpha | + 14\, T_c(|\alpha |) + 50)(t+1)^2\) and \(z = \langle \langle \mathrm{bin}(t), \alpha \rangle , x\rangle \) (with \(\langle \cdot ,\cdot \rangle \) the self-delimiting pair encoding and \(\mathrm{bin}(t)\) the little-endian binary representation of \(t\)): if \(M_\alpha \) on input \(x\) halts within \(t\) steps with output \(w\), then \(U\) on \(z\) halts within \(S\) steps with output \(1w\); and if \(M_\alpha \) does not halt on \(x\) within \(t\) steps, then \(U\) on \(z\) halts within \(S\) steps with output the single bit \(0\).
Let \(c\) be a machine-representation scheme, which assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) (decoding is total). The language \(\mathrm{TMSAT}_c\) consists of all bit strings \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \), where \(\langle a,b\rangle \) is the self-delimiting pair encoding (every bit of \(a\) doubled, then \(01\), then \(b\)) and \(1^k\) is the word of \(k\) ones, for which there exists a certificate \(u\) of length exactly \(n\) such that \(M_\alpha \), started on input \(\langle x,u\rangle \), halts within \(t\) steps with output exactly the single bit \(1\).
For every \(n \in \mathbb {N}\) and every transition action of a one-work-tape binary Turing machine with state set \(\{ 0,\dots ,n\} \), the canonical serialization of the action (the two-bit input-head move, followed by the encodings of the work-tape write, the work-head move, the output emission and the successor state) has length at least \(1\).
Let \(c\) be a machine-representation scheme, which assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) (decoding is total). For every bit string \(y\), \(y \in \mathrm{TMSAT}_c\) if and only if there is a bit string \(w\) with \(|w| = |y| + 1\) and \(yw \in V_c\). Here, with \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \), where \(\langle a,b\rangle \) is the self-delimiting pair encoding and \(1^k\) is the word of \(k\) ones, \(\mathrm{TMSAT}_c\) is the set of strings \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \), over all bit strings \(\alpha , x\) and \(n, t \in \mathbb {N}\), such that \(M_\alpha \) halts on \(\langle x,u\rangle \) within \(t\) steps with output exactly \(1\) for some \(u\) with \(|u| = n\), and \(V_c\) is the set of strings \(y'w'\) with \(|w'| = |y'|+1\) and \(y' = \langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \) for some \(\alpha , x, n, t\) such that \(M_\alpha \) halts on \(\langle x, w'_{[:n]}\rangle \) within \(t\) steps with output exactly \(1\), where \(w'_{[:n]}\) is the first \(n\) bits of \(w'\).
Let \(c\) be an effective machine-representation scheme: it assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \), and comes with a canonizer machine that, on every input \(\alpha \), outputs the canonical serialization of \(M_\alpha \) within \(T_c(|\alpha |)\) steps. For a bit string \(\alpha \), let \(L\) be the length of the canonical serialization of \(M_\alpha \), let \(\{ 0,\dots ,m\} \) be its state set and \(q_0\) its initial state, let \(\mathrm{bin}(m)\) be the binary representation of \(m\), and let \(B = 3L + 5(m+1) + 20\). Then
For a bit string \(z\), write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) holds when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and the width and clock fields consist only of ones. Then the maps sending \(z\) to its code field, input field, width field, clock field and to \(W\), and the map sending \(z\) to the single bit \(G(z)\), are all polynomial-time computable.
Let \(A\) and \(B\) be types and let \(f\) map every \(a \in A\) to a list \(f(a)\) over \(B\) of length at least \(1\). Then for every list \((a_1,\dots ,a_k)\) over \(A\), the concatenation \(f(a_1)\, f(a_2) \cdots f(a_k)\) has length at least \(k\).
For a bit string \(z\), write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) holds when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and the width and clock fields consist only of ones. If \(G(z)\) holds, then the code field and the clock field of \(z\) each have length at most \(|z|\).
Let \(\alpha , x, w\) be bit strings and \(n, t \in \mathbb {N}\), let \(y = \langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \) (\(1^k\) the word of \(k\) ones), suppose \(|w| = |y| + 1\), and let \(z = yw\). Write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) holds when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and the width and clock fields consist only of ones. Then \(G(z)\) holds, the code field of \(z\) is \(\alpha \), its input field is \(x\), its width and clock fields have lengths \(n\) and \(t\), and \(W = w\).
For a bit string \(z\), write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) holds when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and the width and clock fields consist only of ones. If \(G(z)\) holds, then \(z = YW\), \(|W| = |Y| + 1\), and \(Y = \langle \alpha , \langle x, \langle 1^n, 1^t\rangle \rangle \rangle \), where \(\alpha \) and \(x\) are the code and input fields, \(n\) and \(t\) are the lengths of the width and clock fields, and \(1^k\) is the word of \(k\) ones.
For every bit string \(w\), the fixed-width little-endian binary increment of \(w\) (its successor among the words of length \(|w|\), or overflow when there is none) overflows if and only if \(w = 1^{|w|}\), that is, \(w\) consists only of ones.
Consider the fixed-width little-endian binary increment, which sends a bit string \(w\) to its successor among the words of length \(|w|\), or fails (overflow) when \(w\) consists only of ones, the empty word included. For every bit string \(w\), the increment of \(w\) is never the empty word.
Let \(c\) be an effective machine-representation scheme, assigning to each bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) and equipped with a canonizer machine that outputs the canonical serialization of \(M_\alpha \) within \(T_c(|\alpha |)\) steps, and suppose \(T_c\) is polynomially bounded: \(T_c(n) \le K(n+1)^k\) for all \(n\), for some constants \(K, k\). Then \(\mathrm{TMSAT}_c \in \mathrm{NP}\), where \(\mathrm{TMSAT}_c\) is the set of strings \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \), where \(\langle a,b\rangle \) is the self-delimiting pair encoding and \(1^k\) is the word of \(k\) ones, such that some \(u\) with \(|u| = n\) makes \(M_\alpha \) halt on \(\langle x,u\rangle \) within \(t\) steps with output exactly \(1\).
Let \(c\) be an effective machine-representation scheme, assigning to each bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) and equipped with a canonizer machine that outputs the canonical serialization of \(M_\alpha \) within \(T_c(|\alpha |)\) steps, and suppose \(T_c(n) \le K(n+1)^k\) for all \(n\), for some constants \(K, k\). Then \(\mathrm{TMSAT}_c\) is \(\mathrm{NP}\)-complete, where \(\mathrm{TMSAT}_c\) is the set of strings \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \), where \(\langle a,b\rangle \) is the self-delimiting pair encoding and \(1^k\) is the word of \(k\) ones, such that some \(u\) with \(|u| = n\) makes \(M_\alpha \) halt on \(\langle x,u\rangle \) within \(t\) steps with output exactly \(1\).
Let \(c\) be a machine-representation scheme, assigning to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) so that the code of every such machine decodes back to it. Then \(\mathrm{TMSAT}_c\) is \(\mathrm{NP}\)-hard: every language in \(\mathrm{NP}\) reduces to it by a polynomial-time many-one reduction, where \(\mathrm{TMSAT}_c\) is the set of strings \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \), where \(\langle a,b\rangle \) is the self-delimiting pair encoding and \(1^k\) is the word of \(k\) ones, such that some \(u\) with \(|u| = n\) makes \(M_\alpha \) halt on \(\langle x,u\rangle \) within \(t\) steps with output exactly \(1\).
For all bit strings \(\alpha , x\) and all \(n, t \in \mathbb {N}\), each of \(|\alpha |\), \(|x|\), \(n\) and \(t\) is at most the length of the quadruple encoding \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \), where \(\langle a,b\rangle \) is the self-delimiting pair encoding and \(1^k\) is the word of \(k\) ones.
For bit strings \(\alpha , x, \beta , y\) and \(n, t, m, s \in \mathbb {N}\), if \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle = \langle \beta ,\langle y,\langle 1^m,1^s\rangle \rangle \rangle \), where \(\langle a,b\rangle \) is the self-delimiting pair encoding and \(1^k\) the word of \(k\) ones, then \(\alpha = \beta \), \(x = y\), \(n = m\) and \(t = s\).
Let \(c\) be a machine-representation scheme, which assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) (decoding is total). Let \(L\) and \(V\) be languages of bit strings, \(C, e \in \mathbb {N}\), \(\alpha \) a bit string and \(T \colon \mathbb {N} \to \mathbb {N}\). Suppose that \(x \in L\) if and only if some \(u\) with \(|u| = C(|x|+1)^e\) satisfies \(xu \in V\), and that for all \(x, u\) with \(|u| = C(|x|+1)^e\), \(M_\alpha \) on input \(\langle x,u\rangle \) halts within \(T(|x|)\) steps with output the single bit \([xu \in V]\). Then for every bit string \(x\), \(x \in L\) if and only if \(\langle \alpha ,\langle x,\langle 1^{C(|x|+1)^e},1^{T(|x|)}\rangle \rangle \rangle \in \mathrm{TMSAT}_c\), where \(\langle a,b\rangle \) is the self-delimiting pair encoding, \(1^k\) the word of \(k\) ones, and \(\mathrm{TMSAT}_c\) the set of strings \(\langle \beta ,\langle y,\langle 1^n,1^t\rangle \rangle \rangle \) such that \(M_\beta \) halts on \(\langle y,v\rangle \) within \(t\) steps with output exactly \(1\) for some \(v\) with \(|v| = n\).
For a bit string \(z\), write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) holds when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and the width and clock fields consist only of ones, and the timed request of \(z\) is \(\langle \langle \mathrm{bin}(t), \alpha \rangle , \langle x, W_{[:n]}\rangle \rangle \) when \(G(z)\) holds, where \(\alpha \) and \(x\) are the code and input fields, \(n\) and \(t\) the lengths of the width and clock fields, \(W_{[:n]}\) the first \(n\) bits of \(W\), and \(\mathrm{bin}(t)\) the little-endian binary representation of \(t\); otherwise it is the fixed request \(\langle \langle \varepsilon ,\varepsilon \rangle ,\varepsilon \rangle \). The map sending \(z\) to its timed request is polynomial-time computable.
Let \(c\) be a machine-representation scheme, which assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) (decoding is total). For a bit string \(z\), write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) holds when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and the width and clock fields consist only of ones, and the result for \(z\) is \(A_c(\alpha , \langle x, W_{[:n]}\rangle , t)\) when \(G(z)\) holds, with \(\alpha \), \(x\) the code and input fields, \(n\), \(t\) the lengths of the width and clock fields and \(W_{[:n]}\) the first \(n\) bits of \(W\), and \(A_c(\varepsilon ,\varepsilon ,0)\) otherwise; here \(A_c(\beta , v, t')\) is \(1w'\) if \(M_\beta \) run on \(v\) for \(t'\) steps has halted with output \(w'\), and the single bit \(0\) if it has not halted. Then for every bit string \(z\), the result for \(z\) equals \(11\) if and only if \(z \in V_c\), where \(V_c\) is the set of strings \(y'w'\) with \(|w'| = |y'| + 1\) for which there are bit strings \(\beta , v\) and \(n', t' \in \mathbb {N}\) with \(y' = \langle \beta ,\langle v,\langle 1^{n'},1^{t'}\rangle \rangle \rangle \) such that \(M_\beta \) halts on \(\langle v, w'_{[:n']}\rangle \) within \(t'\) steps with output exactly \(1\), \(w'_{[:n']}\) being the first \(n'\) bits of \(w'\).
Let \(c\) be an effective machine-representation scheme: it assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \), and comes with a canonizer machine that, on every input \(\alpha \), outputs the canonical serialization of \(M_\alpha \) within \(T_c(|\alpha |)\) steps. Then for every bit string \(\alpha \), the canonical serialization of \(M_\alpha \) has length at most \(T_c(|\alpha |)\).
Let \(M\) be a one-work-tape binary Turing machine in code normal form with state set \(\{ 0,\dots ,m\} \) and initial state \(q_0\), and let \(L\) be the length of its canonical serialization. Then the binary representation of \(m\) has length at most \(L\), \(q_0 \le L\), and \(m + 1 \le L\).
Let \(c\) be an effective machine-representation scheme: it assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \), and comes with a canonizer machine that, on every input \(\alpha \), outputs the canonical serialization of \(M_\alpha \) within \(T_c(|\alpha |)\) steps, and suppose \(T_c\) is polynomially bounded: \(T_c(n) \le K(n+1)^k\) for all \(n\), for some constants \(K, k\). Then there exist \(A, d \in \mathbb {N}\) such that for all \(m, r, t \in \mathbb {N}\) with \(r \le m\) and \(t \le m\),
Let \(c\) be an effective machine-representation scheme: it assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \), and comes with a canonizer machine that, on every input \(\alpha \), outputs the canonical serialization of \(M_\alpha \) within \(T_c(|\alpha |)\) steps. For bit strings \(\alpha , x\) and \(t \in \mathbb {N}\) write \(S = (3|\alpha |+14\, T_c(|\alpha |)+50)(t+1)^2\) and \(z = \langle \langle \mathrm{bin}(t),\alpha \rangle ,x\rangle \) (self-delimiting pair encoding, little-endian binary \(\mathrm{bin}(t)\)). Let \(U\) be a finite binary Turing machine such that, for all \(\alpha , x, t\), whenever \(M_\alpha \) on \(x\) halts within \(t\) steps with output \(w\), \(U\) on \(z\) halts within \(S\) steps with output \(1w\), and whenever \(M_\alpha \) does not halt on \(x\) within \(t\) steps, \(U\) on \(z\) halts within \(S\) steps with output \(0\). Then for all \(\alpha , x, t\), \(U\) on \(z\) halts within \(S\) steps with output exactly \(A_c(\alpha ,x,t)\), which is \(1w\) if \(M_\alpha \) run on \(x\) for \(t\) steps has halted with output \(w\), and the single bit \(0\) otherwise.
Let \(z\) be a bit string such that \(s(z)\) parses as a pair encoding, where \(s(z)\) is the odd split of \(z\) (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) of the first \(i\) bits of \(z\) and the rest when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), and \(\pi _1, \pi _2\) return the first and second components of a pair encoding (the empty word on strings that do not parse as pairs). Then \(z = \pi _1(s(z))\, \pi _2(s(z))\) and \(|\pi _2(s(z))| = |\pi _1(s(z))| + 1\).
For all bit strings \(a\) and \(b\), the prefix-extraction machine \(E\) (one work tape, four states; it counts the data blocks of a pair encoding in unary and then copies that many payload bits), started on the self-delimiting pair encoding \(\langle a,b\rangle \), halts within \(|\langle a,b\rangle | + 1\) steps with output exactly the first \(|a|\) bits of \(b\) (all of \(b\) if \(|b| \lt |a|\)).
Let \(x\) be a bit string, \(i, m \in \mathbb {N}\) with \(i + 2 \le |x|\), and \(b\) a bit with \(x_i = x_{i+1} = b\) (positions counted from \(0\)). Then running \(E\) for \(2\) steps from \(\Delta _x(0, i, m, m, \varepsilon )\) yields \(\Delta _x(0, i+2, m+1, m+1, \varepsilon )\); here \(E\) is the prefix-extraction machine (one work tape, states \(\{ 0,1,2,3\} \)) that counts the aligned data blocks \(00\) and \(11\) of a pair encoding \(\langle a,b\rangle \) in unary on its work tape and then copies bits of \(b\) to its output while moving its work head back along the counter, and \(\Delta _x(q,i,m,h,o)\) denotes the configuration of \(E\) on input \(x\) in state \(q\), with the input head on input position \(i\) (counting from \(0\)), the work tape holding \(1^m\) in cells \(0,\dots ,m-1\) and blanks elsewhere, the work head at \(h\), and output \(o\).
Let \(a, b, p\) be bit strings, \(x = p\, \langle a,b\rangle \) with \(\langle a,b\rangle \) the self-delimiting pair encoding, and \(m \in \mathbb {N}\). Then running \(E\) for \(2|a| + 2\) steps from \(\Delta _x(0, |p|, m, m, \varepsilon )\) yields \(\Delta _x(3, |p| + 2|a| + 2, m + |a|, m + |a| - 1, \varepsilon )\); here \(E\) is the prefix-extraction machine (one work tape, states \(\{ 0,1,2,3\} \)) that counts the aligned data blocks \(00\) and \(11\) of a pair encoding \(\langle a,b\rangle \) in unary on its work tape and then copies bits of \(b\) to its output while moving its work head back along the counter, and \(\Delta _x(q,i,m,h,o)\) denotes the configuration of \(E\) on input \(x\) in state \(q\), with the input head on input position \(i\) (counting from \(0\)), the work tape holding \(1^m\) in cells \(0,\dots ,m-1\) and blanks elsewhere, the work head at \(h\), and output \(o\).
Let \(p\) and \(b\) be bit strings, \(x = pb\), \(m \in \mathbb {N}\), and \(j \le \min (m, |b|)\). Then running \(E\) for \(j\) steps from \(\Delta _x(3, |p|, m, m-1, \varepsilon )\) yields \(\Delta _x(3, |p| + j, m, m - j - 1, b_{[:j]})\), where \(b_{[:j]}\) is the first \(j\) bits of \(b\); here \(E\) is the prefix-extraction machine (one work tape, states \(\{ 0,1,2,3\} \)) that counts the aligned data blocks \(00\) and \(11\) of a pair encoding \(\langle a,b\rangle \) in unary on its work tape and then copies bits of \(b\) to its output while moving its work head back along the counter, and \(\Delta _x(q,i,m,h,o)\) denotes the configuration of \(E\) on input \(x\) in state \(q\), with the input head on input position \(i\) (counting from \(0\)), the work tape holding \(1^m\) in cells \(0,\dots ,m-1\) and blanks elsewhere, the work head at \(h\), and output \(o\).
Let \(x\) be a bit string, \(q\) an optional state in \(\{ 0,1,2,3\} \), \(i \le |x|\), \(m \in \mathbb {N}\), \(h \in \mathbb {Z}\) and \(o\) an output word. In the configuration of a one-work-tape binary machine with states \(\{ 0,1,2,3\} \) on input \(x\) that is in state \(q\), has its input head on input position \(i\) (counting from \(0\)), its work tape holding \(1^m\) in cells \(0,\dots ,m-1\), its work head at \(h\) and output \(o\), the scanned input symbol is \(x_i\) if \(i \lt |x|\) and the blank if \(i = |x|\).
Let \(x\) be a bit string and \(i, m \in \mathbb {N}\) with \(i + 2 \le |x|\), \(x_i = 0\) and \(x_{i+1} = 1\) (positions counted from \(0\)). Then running \(E\) for \(2\) steps from \(\Delta _x(0, i, m, m, \varepsilon )\) yields \(\Delta _x(3, i+2, m, m-1, \varepsilon )\); here \(E\) is the prefix-extraction machine (one work tape, states \(\{ 0,1,2,3\} \)) that counts the aligned data blocks \(00\) and \(11\) of a pair encoding \(\langle a,b\rangle \) in unary on its work tape and then copies bits of \(b\) to its output while moving its work head back along the counter, and \(\Delta _x(q,i,m,h,o)\) denotes the configuration of \(E\) on input \(x\) in state \(q\), with the input head on input position \(i\) (counting from \(0\)), the work tape holding \(1^m\) in cells \(0,\dots ,m-1\) and blanks elsewhere, the work head at \(h\), and output \(o\).
Let \(c\) be a machine-representation scheme, which assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) (decoding is total). For bit strings \(\alpha , x\) and \(t \in \mathbb {N}\), the answer word \(A_c(\alpha , x, t)\) is \(1w\) if \(M_\alpha \), run for \(t\) steps on input \(x\), has halted with output \(w\), and is the single bit \(0\) if it has not halted after \(t\) steps.
Let \(c\) be a machine-representation scheme, which assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) (decoding is total). For all bit strings \(\alpha , x\) and \(t \in \mathbb {N}\), let \(A_c(\alpha ,x,t)\) be \(1w\) if \(M_\alpha \) run on \(x\) for \(t\) steps has halted with output \(w\), and the single bit \(0\) otherwise. Then \(A_c(\alpha ,x,t) = 11\) if and only if \(M_\alpha \) on input \(x\) halts within \(t\) steps with output exactly the single bit \(1\).
For a bit string \(z\), the clock field is \(\pi _2(\pi _2(\pi _2(Y(z))))\), where \(Y(z) = \pi _1(s(z))\) and \(s(z)\) is the odd split of \(z\) (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) of the first \(i\) bits of \(z\) and the rest when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), and \(\pi _1, \pi _2\) return the first and second components of a pair encoding (the empty word on strings that do not parse as pairs); on well-formed inputs it is the unary deadline \(1^t\), but its shape is not checked.
For a bit string \(z\), the code field is \(\pi _1(Y(z))\), where \(Y(z) = \pi _1(s(z))\) and \(s(z)\) is the odd split of \(z\) (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) of the first \(i\) bits of \(z\) and the rest when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), and \(\pi _1, \pi _2\) return the first and second components of a pair encoding (the empty word on strings that do not parse as pairs).
This is the function on bit strings that sends every valid self-delimiting pair encoding \(\langle a,b\rangle \) (the bits of \(a\) doubled, then \(01\), then \(b\)) to the concatenation \(ab\), and every bit string that does not parse as a pair encoding to the empty string.
For a bit string \(z\), write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) is the Boolean that is true exactly when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and both the width field and the clock field consist only of ones (the empty word included).
For a bit string \(z\), the input field is \(\pi _1(\pi _2(Y(z)))\), where \(Y(z) = \pi _1(s(z))\) and \(s(z)\) is the odd split of \(z\) (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) of the first \(i\) bits of \(z\) and the rest when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), and \(\pi _1, \pi _2\) return the first and second components of a pair encoding (the empty word on strings that do not parse as pairs).
For bit strings \(\alpha , x\) and \(n, t \in \mathbb {N}\), the quadruple encoding \(\langle \alpha , x, 1^n, 1^t\rangle \) is the bit string \(\langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \), where \(\langle a,b\rangle \) is the self-delimiting pair encoding and \(1^k\) is the word of \(k\) ones.
For a bit string \(z\), write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) holds when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and the width and clock fields consist only of ones. Then the timed request of \(z\) is \(\langle \langle \mathrm{bin}(t), \alpha \rangle , \langle x, W_{[:n]}\rangle \rangle \) when \(G(z)\) holds, where \(\alpha \) and \(x\) are the code and input fields, \(n\) and \(t\) the lengths of the width and clock fields, \(W_{[:n]}\) the first \(n\) bits of \(W\), and \(\mathrm{bin}(t)\) the little-endian binary representation of \(t\); otherwise it is the fixed request \(\langle \langle \varepsilon ,\varepsilon \rangle ,\varepsilon \rangle \).
Let \(c\) be a machine-representation scheme, which assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) (decoding is total). For a bit string \(z\), write \(s(z)\) for its odd split (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), \(\pi _1, \pi _2\) for the two components of a pair encoding (the empty word when parsing fails), \(Y = \pi _1(s(z))\) and \(W = \pi _2(s(z))\); the code, input, width and clock fields of \(z\) are \(\pi _1(Y)\), \(\pi _1(\pi _2(Y))\), \(\pi _1(\pi _2(\pi _2(Y)))\) and \(\pi _2(\pi _2(\pi _2(Y)))\). The guard \(G(z)\) holds when \(s(z)\), \(Y\), \(\pi _2(Y)\) and \(\pi _2(\pi _2(Y))\) all parse as pair encodings and the width and clock fields consist only of ones. Then the result for \(z\) is \(A_c(\alpha , \langle x, W_{[:n]}\rangle , t)\) when \(G(z)\) holds, with \(\alpha \), \(x\) the code and input fields, \(n\), \(t\) the lengths of the width and clock fields and \(W_{[:n]}\) the first \(n\) bits of \(W\), and \(A_c(\varepsilon ,\varepsilon ,0)\) otherwise; here \(A_c(\beta , v, t')\) is \(1w'\) if \(M_\beta \) run on \(v\) for \(t'\) steps has halted with output \(w'\), and the single bit \(0\) if it has not halted.
For a bit string \(z\), the odd split \(s(z)\) is the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) of the first \(i\) bits of \(z\) and the remaining bits, where \(i\) is the least index in \(\{ 0,\dots ,|z|\} \) with \(i + (i+1) = |z|\); if no such index exists (that is, \(|z|\) is even), \(s(z)\) is the empty string.
For a bit string \(x\), an optional state \(q \in \{ 0,1,2,3\} \) (absence meaning halted), an index \(i \le |x|\), \(m \in \mathbb {N}\), a work-head position \(h \in \mathbb {Z}\) and an output word \(o\), \(\Delta _x(q,i,m,h,o)\) is the configuration of a one-work-tape binary machine with states \(\{ 0,1,2,3\} \) on input \(x\) in state \(q\), with the input head on input position \(i\) (counting from \(0\); the blank cell right of the input when \(i = |x|\)), the work tape holding \(1^m\) in cells \(0,\dots ,m-1\) and blanks elsewhere, the work head at \(h\), and output \(o\).
The prefix extractor \(E\) is the finite binary-alphabet machine with one work tape, states \(\{ 0,1,2,3\} \) and initial state \(0\) which, on an input \(\langle a,b\rangle \) (the bits of \(a\) doubled, then \(01\), then \(b\)), reads the doubled region in aligned blocks of two bits: each data block \(00\) or \(11\) writes one more \(1\) on a unary counter on its work tape and moves the work head right, and the separator \(01\) moves the work head left onto the last counter cell and enters state \(3\). In state \(3\) it copies the next input bit to the output and moves its input head right and its work head left, halting as soon as the work head reads a blank or the input is exhausted; a malformed block \(10\), or the end of the input in states \(0\) to \(2\), halts it immediately without output.
Let \(c\) be a machine-representation scheme, which assigns to every bit string \(\alpha \) a one-work-tape binary Turing machine \(M_\alpha \) (decoding is total). The verifier language \(V_c\) consists of all bit strings \(yw\) such that \(|w| = |y| + 1\), \(y = \langle \alpha ,\langle x,\langle 1^n,1^t\rangle \rangle \rangle \) for some bit strings \(\alpha , x\) and some \(n, t \in \mathbb {N}\) (self-delimiting pair encoding, \(1^k\) the word of \(k\) ones), and \(M_\alpha \) on input \(\langle x, w_{[:n]}\rangle \) halts within \(t\) steps with output exactly \(1\), where \(w_{[:n]}\) consists of the first \(n\) bits of \(w\).
For a bit string \(z\), the padded certificate part is \(W(z) = \pi _2(s(z))\), where \(s(z)\) is the odd split of \(z\) (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) of the first \(i\) bits of \(z\) and the rest when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), and \(\pi _1, \pi _2\) return the first and second components of a pair encoding (the empty word on strings that do not parse as pairs).
For a bit string \(z\), the width field is \(\pi _1(\pi _2(\pi _2(Y(z))))\), where \(Y(z) = \pi _1(s(z))\) and \(s(z)\) is the odd split of \(z\) (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) of the first \(i\) bits of \(z\) and the rest when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), and \(\pi _1, \pi _2\) return the first and second components of a pair encoding (the empty word on strings that do not parse as pairs); on well-formed inputs it is the unary certificate length \(1^n\), but its shape is not checked.
For a language \(V\) of bit strings, this is the total function on bit strings that sends a valid self-delimiting pair encoding \(\langle x,u\rangle \) to the single bit \(1\) if the concatenation \(xu\) lies in \(V\) and to \(0\) otherwise, and sends every string that does not parse as a pair encoding to the single bit \(0\).
For a bit string \(z\), the instance part is \(Y(z) = \pi _1(s(z))\), where \(s(z)\) is the odd split of \(z\) (the self-delimiting pair encoding \(\langle z_{[:i]}, z_{[i:]}\rangle \) of the first \(i\) bits of \(z\) and the rest when \(|z| = 2i+1\), and the empty word when \(|z|\) is even), and \(\pi _1, \pi _2\) return the first and second components of a pair encoding (the empty word on strings that do not parse as pairs).
Let \(\sigma \) be a set of control states, let \(\delta : \sigma \times \{ 0,1\} \to \sigma \) be a transition function, and let \(o : \sigma \times \{ 0,1\} \to \{ 0,1\} \cup \{ \bot \} \) be an emission function, where the value \(\bot \) means that nothing is emitted. The string function of the one-pass transducer \((\delta , o)\) started in state \(q \in \sigma \) is the map \(w \mapsto \tau _{\delta ,o}(q, w)\) on \(\{ 0,1\} ^*\) defined recursively by \(\tau _{\delta ,o}(q, \varepsilon ) = \varepsilon \) and
for a bit \(b\) and a string \(w\), where \(\bar o(q,b)\) is the one-bit word \(o(q,b)\) when \(o(q,b) \in \{ 0,1\} \) and the empty word when \(o(q,b) = \bot \).
Let \(\sigma \) be a set of states, \(\delta : \sigma \times \{ 0,1\} \to \sigma \) a transition function and \(o : \sigma \times \{ 0,1\} \to \{ 0,1\} \cup \{ \bot \} \) an emission function, and write \(\tau _{\delta ,o}(q, w)\) for the output of the one-pass transducer that reads \(w\) from left to right starting in state \(q\), where each bit \(b\) read in state \(p\) emits the bit \(o(p,b)\) (or nothing if \(o(p,b) = \bot \)) and moves to state \(\delta (p,b)\). Then for every state \(q\) and all bit strings \(u, v \in \{ 0,1\} ^*\),
where \(\hat\delta (q,u)\) is the state reached from \(q\) by applying \(\delta \) successively to the bits of \(u\) from left to right.
Let \(\sigma \) be a finite set of states, \(\delta : \sigma \times \{ 0,1\} \to \sigma \), \(o : \sigma \times \{ 0,1\} \to \{ 0,1\} \cup \{ \bot \} \) and \(q_0 \in \sigma \), and let \(M\) be the binary Turing machine with no work tapes, state set \(\sigma \) and start state \(q_0\) that, in state \(p\) reading an input bit \(b\), appends \(o(p,b)\) to its output (nothing if \(o(p,b) = \bot \)), moves its input head one cell right and enters \(\delta (p,b)\), and that halts on reading a blank. Let \(x = x_0 x_1 \cdots x_{n-1} \in \{ 0,1\} ^*\), let \(0 \le i \le n\), let \(w = x_i x_{i+1} \cdots x_{n-1}\) be the suffix of \(x\) starting at position \(i\), and let \(C\) be any configuration of \(M\) on input \(x\) that is in state \(q \in \sigma \) with its input head on the cell holding \(x_i\) (on the blank just right of the input when \(i = n\)). Then running \(M\) for \(|w| + 1\) steps from \(C\) reaches the halting state, and the output tape then holds the output of \(C\) followed by \(\tau _{\delta ,o}(q, w)\), where \(\tau _{\delta ,o}(q, w)\) denotes the word emitted by reading \(w\) from left to right starting in state \(q\), each bit \(b\) read in state \(p\) emitting \(o(p,b)\) (or nothing) and moving to \(\delta (p,b)\).
Let \(\sigma \) be a finite set of states, \(\delta : \sigma \times \{ 0,1\} \to \sigma \) a transition function, \(o : \sigma \times \{ 0,1\} \to \{ 0,1\} \cup \{ \bot \} \) an emission function (\(\bot \) meaning no output) and \(q_0 \in \sigma \) a start state. The transducer machine \(M_{\delta ,o,q_0}\) is the finite binary Turing machine with no work tapes, state set \(\sigma \) and start state \(q_0\) whose transition function, in state \(q\) reading an input bit \(b\), appends \(o(q,b)\) to the output tape (nothing if \(\bot \)), moves the input head one cell right and enters state \(\delta (q,b)\), and on reading a blank input cell halts without moving or writing.
Let \(\sigma \) be a finite set of states, \(\delta : \sigma \times \{ 0,1\} \to \sigma \) a transition function, \(o : \sigma \times \{ 0,1\} \to \{ 0,1\} \cup \{ \bot \} \) an emission function (\(\bot \) meaning no output) and \(q_0 \in \sigma \), and let \(M\) be the binary Turing machine with no work tapes, state set \(\sigma \) and start state \(q_0\) that, in state \(p\) reading an input bit \(b\), appends \(o(p,b)\) to its output (nothing if \(\bot \)), moves its input head right and enters \(\delta (p,b)\), and halts on reading a blank. Then \(M\) computes the function \(x \mapsto \tau _{\delta ,o}(q_0, x)\) in time \(n + 1\): on every input \(x \in \{ 0,1\} ^*\) it halts within \(|x| + 1\) steps with \(\tau _{\delta ,o}(q_0, x)\) on its output tape, where \(\tau _{\delta ,o}(q_0, x)\) is the word emitted by reading \(x\) from left to right starting in \(q_0\), each bit \(b\) read in state \(p\) emitting \(o(p,b)\) (or nothing) and moving to \(\delta (p,b)\).
Let \(\sigma \) be a set of states, \(\delta : \sigma \times \{ 0,1\} \to \sigma \) a transition function and \(o : \sigma \times \{ 0,1\} \to \{ 0,1\} \cup \{ \bot \} \) an emission function (\(\bot \) meaning no output). This defines the transition function, for a binary Turing machine with state set \(\sigma \), a read-only input tape, no work tapes and a write-only output tape, given as follows. In state \(q\) reading an input bit \(b\), the machine appends \(o(q,b)\) to the output tape (nothing if \(o(q,b) = \bot \)), moves its input head one cell to the right and enters state \(\delta (q,b)\); in any state, reading a blank input cell, it leaves the head in place, writes nothing and halts.
Fix a representation scheme \(c\) for Turing machines in code normal form (one work tape, binary alphabet): a total decoding assigning to every bit string a machine, together with an encoding that the decoding inverts. The diagonal function \(\mathrm{UC}_c : \{ 0,1\} ^* \to \{ 0,1\} \) maps a bit string \(\alpha \) to \(\mathrm{false}\) if the machine that \(\alpha \) denotes under \(c\), run on the input \(\alpha \) itself, halts after some finite number of steps with completed output the one-bit string \([\mathrm{true}]\) — that is, if the machine accepts its own representation. In every other case — the machine diverges on \(\alpha \), or halts with any completed output other than \([\mathrm{true}]\) — the value is \(\mathrm{true}\). ([AB09, §1.5] writes: \(UC(\alpha ) = 0\) if \(M_\alpha (\alpha ) = 1\), and \(UC(\alpha ) = 1\) otherwise.)
Let \(c\) be an effective coding scheme for Turing machines, i.e. one admitting a universal evaluator machine. Write \(\mathrm{HALT}_c\) for the halting function of \(c\): it is true on a bit string \(s\) exactly when \(s\) is the pairing encoding \(\langle \alpha , x \rangle \) of two bit strings and the machine denoted by \(\alpha \) under \(c\) halts on input \(x\), and false otherwise. Write \(\mathrm{UC}_c\) for the diagonal function that maps a bit string \(\alpha \) to false when the machine denoted by \(\alpha \) under \(c\) outputs the single bit true on input \(\alpha \) within some finite number of steps, and to true otherwise. If the function mapping \(s\) to the one-bit string carrying \(\mathrm{HALT}_c(s)\) is computable by a Turing machine, then so is the function mapping \(\alpha \) to the one-bit string carrying \(\mathrm{UC}_c(\alpha )\) [AB09, proof of Theorem 1.11].
Let \(c\) be a representation scheme for Turing machines in code normal form (one work tape, binary alphabet; a total decoding of bit strings to machines, inverted by the encoding), and let \(\alpha \) be a bit string. Then the diagonal function of \(c\) takes the value \(\mathrm{false}\) at \(\alpha \) — i.e. \(\mathrm{UC}_c(\alpha ) = \mathrm{false}\) — if and only if there exists a number of steps \(t\) within which the machine denoted by \(\alpha \) under \(c\), run on the input \(\alpha \) itself, halts with completed output the one-bit string \([\mathrm{true}]\).
Let \(c\) be a representation scheme for Turing machines in code normal form (one work tape, binary alphabet; a total decoding of bit strings to machines, inverted by the encoding), and let \(\alpha \) be a bit string. Then the diagonal function of \(c\) takes the value \(\mathrm{true}\) at \(\alpha \) — i.e. \(\mathrm{UC}_c(\alpha ) = \mathrm{true}\) — if and only if there is no number of steps \(t\) within which the machine denoted by \(\alpha \) under \(c\), run on the input \(\alpha \) itself, halts with completed output the one-bit string \([\mathrm{true}]\): the machine diverges on \(\alpha \), or halts with a completed output other than \([\mathrm{true}]\).
Let \(c\) be any representation scheme for Turing machines in code normal form (one work tape, binary alphabet; a total decoding of bit strings to machines, inverted by the encoding), effective or not, and let \(\mathrm{UC}_c\) be its diagonal function: \(\mathrm{UC}_c(\alpha ) = \mathrm{false}\) if the machine denoted by \(\alpha \) under \(c\) halts on the input \(\alpha \) itself with completed output the one-bit string \([\mathrm{true}]\), and \(\mathrm{UC}_c(\alpha ) = \mathrm{true}\) otherwise. Then the total function on bit strings \(\alpha \mapsto [\mathrm{UC}_c(\alpha )]\), returning the one-bit string whose single bit is the value of the diagonal function, is not computable: no finite Turing machine over the binary alphabet halts on every input \(\alpha \) with \([\mathrm{UC}_c(\alpha )]\) on its output tape [AB09, Theorem 1.10].
The controller phase used to rewind the unary counter: the unary-reset phase when the Boolean flag marks the final rewind, and the unary-back phase otherwise. Both rewinds follow the same path and differ only in their successor phase.
Let \(L_1, L_2 \subseteq \{ 0,1\} ^*\) be languages in \(\mathsf{P}\), the class of languages decidable by a deterministic Turing machine in polynomial time. Then their union \(L_1 \cup L_2 = \{ z : z \in L_1 \text{ or } z \in L_2\} \) also belongs to \(\mathsf{P}\).
For \(C, c \in \mathbb {N}\), the Boolean function \(\beta _{C,c}\) on binary strings that, on a string \(z\) which the aligned pair parser decodes as a pair \((a, b)\) — that is, \(z = \langle a, b\rangle \) is a well-formed self-delimiting pair encoding (\(a\) with every bit doubled, then \(01\), then \(b\)) — returns true exactly when \(|b| \le C(|a|+1)^c\), and returns false on every string that is not a well-formed pair encoding.
The function on binary strings that sends \(z\) to the concatenation \(ab\) when the aligned pair parser decodes \(z\) as a pair \((a, b)\) — that is, when \(z = \langle a, b\rangle \) is a well-formed self-delimiting pair encoding (\(a\) with every bit doubled, then \(01\), then \(b\)) — and sends every other string to the empty string.
Let \(V \subseteq \{ 0,1\} ^{*}\) be a language and let \(\mathbf{1}_V(x)\) be the bit \(1\) if \(x \in V\) and \(0\) otherwise. If the map sending a binary string \(x\) to the one-bit string \(\mathbf{1}_V(x)\) is polynomial-time computable, i.e. some finite multi-tape Turing machine over the binary alphabet computes it within \(D(n+1)^d\) steps on every input of length \(n\) for some constants \(D, d\), then \(V \in \mathsf{P}\).
For \(C, c \in \mathbb {N}\) let \(\beta _{C,c}(z)\), for a binary string \(z\), be the bit \(1\) if \(z = \langle a, b\rangle \) is a well-formed self-delimiting pair encoding (\(a\) with every bit doubled, then \(01\), then \(b\)) with \(|b| \le C(|a|+1)^c\), and the bit \(0\) otherwise. Then for all \(C, c \in \mathbb {N}\) the map sending \(z\) to the one-bit string \(\beta _{C,c}(z)\) is polynomial-time computable: some finite multi-tape Turing machine over the binary alphabet computes it within \(D(n+1)^d\) steps on every input of length \(n\), for some constants \(D, d \in \mathbb {N}\).
Let \(V \subseteq \{ 0,1\} ^{*}\) be a language in \(\mathsf{P}\), and let \(\mathbf{1}_V(x)\) be the bit \(1\) if \(x \in V\) and \(0\) otherwise. Then the map sending a binary string \(x\) to the one-bit string \(\mathbf{1}_V(x)\) is polynomial-time computable: some finite multi-tape Turing machine over the binary alphabet computes it within \(D(n+1)^d\) steps on every input of length \(n\), for some constants \(D, d \in \mathbb {N}\).
For a binary string \(z\), let \(\pi _1(z)\) and \(\pi _2(z)\) be the two components \(a\) and \(b\) when \(z = \langle a, b\rangle \) is a well-formed self-delimiting pair encoding (\(a\) with every bit doubled, then \(01\), then \(b\)), and let both be the empty string otherwise. Then for all \(C, c \in \mathbb {N}\), the map sending \(z\) to the one-bit string recording whether \(C(|\pi _1(z)|+1)^c \le |\pi _2(z)|\) is polynomial-time computable: some finite multi-tape Turing machine over the binary alphabet computes it within \(D(n+1)^d\) steps on every input of length \(n\), for some constants \(D, d \in \mathbb {N}\).
For \(C, c \in \mathbb {N}\), the function on binary strings that, given \(y\), finds the least \(n \le |y|\) with \(n + (C+1)(n+1)^c = |y|\); if such an \(n\) exists it outputs the self-delimiting pair encoding \(\langle x, v\rangle \) (\(x\) with every bit doubled, then \(01\), then \(v\)), where \(x\) is the first \(n\) bits of \(y\) and \(v\) the remaining bits, and otherwise it outputs the empty string.
For \(D, e, m \in \mathbb {N}\) let \(\mathrm{solve}_{D,e}(m)\) be the least \(i \le m\) with \(i + D(i+1)^e = m\), undefined if there is none, and for \(C, c, m \in \mathbb {N}\) let \(\mathrm{split}_{C,c}(m)\) be the least \(n \le m\) with \(n + (C+1)(n+1)^c = m\), undefined if there is none. Then for all \(C, c \in \mathbb {N}\), the partial functions \(\mathrm{solve}_{C+1,c}\) and \(\mathrm{split}_{C,c}\) on \(\mathbb {N}\) are equal.
The function on binary strings that, on a well-formed self-delimiting pair encoding \(z = \langle a, v\rangle \) (\(a\) with every bit doubled, then \(01\), then \(v\)) in which \(v\) contains a \(1\), outputs \(\langle a, u\rangle \), where \(u\) is the prefix of \(v\) preceding its last \(1\); on every other string (a malformed pair, or a pair whose second component has no \(1\)) it outputs the empty string.
Let \(\mathrm{strip}\) be the partial map on binary strings defined by: \(\mathrm{strip}(\varepsilon )\) is undefined, and for a bit \(b\) and a string \(v\), \(\mathrm{strip}(bv) = b\, \mathrm{strip}(v)\) when \(\mathrm{strip}(v)\) is defined, while otherwise \(\mathrm{strip}(1v) = \varepsilon \) and \(\mathrm{strip}(0v)\) is undefined. Let \(\mathrm{last}(v)\) be computed by reversing \(v\), discarding the leading \(0\)s of the reversal, and, if a \(1\) remains, deleting it and returning the reversal of what follows it, while \(\mathrm{last}(v)\) is undefined if no \(1\) remains. Then \(\mathrm{last}\) and \(\mathrm{strip}\) are equal as partial maps on \(\{ 0,1\} ^{*}\).
Let \(M\) be a binary multi-tape Turing machine with \(k\) work tapes whose cells are indexed by \(\mathbb {Z}\), let \(x \in \{ 0,1\} ^*\), \(\tau \in \{ 0, \dots , k-1\} \), \(p \in \mathbb {Z}\), and let \(a \le b\) be natural numbers. If for every \(r\) with \(a \le r \lt b\) the head of work tape \(\tau \) is not at cell \(p\) after \(r\) steps of \(M\) on \(x\), then the content of cell \(p\) of work tape \(\tau \) after \(b\) steps of \(M\) on \(x\) equals its content after \(a\) steps.
Let \(M\) be a binary multi-tape Turing machine with finite state set \(Q\), transition function \(\delta \) and \(k\) work tapes whose cells are indexed by \(\mathbb {Z}\); let \(x \in \{ 0,1\} ^*\), \(t \in \mathbb {N}\), \(\tau \in \{ 0, \dots , k-1\} \) and \(p \in \mathbb {Z}\). Let \(q \in Q \cup \{ \bot \} \) (with \(\bot \) meaning halted), \(a\) and \(b = (b_0, \dots , b_{k-1})\) be the state, the input-head symbol and the work-head symbols after \(t\) steps of \(M\) on \(x\), and let \(h\) be the position of the head of work tape \(\tau \) at that time. Then after \(t + 1\) steps, cell \(p\) of work tape \(\tau \) has the same content as after \(t\) steps if \(p \ne h\); if \(p = h\), its content is \(b_\tau \) when \(q = \bot \), and otherwise is the symbol that \(\delta (q, a, b)\) writes on tape \(\tau \), or \(b_\tau \) when that transition writes nothing on tape \(\tau \).
Let \(M\) be a binary multi-tape Turing machine with \(k\) work tapes whose cells are indexed by \(\mathbb {Z}\), and let \(n, t \in \mathbb {N}\) and \(\tau \in \{ 0, \dots , k-1\} \). The scheduled work position \(\mathrm{wpos}_M(n, t, \tau ) \in \mathbb {Z}\) is the position of the head of work tape \(\tau \) after \(t\) steps of \(M\) on the all-zeros input \(0^n\).
Let \(M\) be a binary multi-tape Turing machine with finite state set \(Q\), \(k\) work tapes and transition function \(\delta \), let \((q, a, b)\) be a snapshot of \(M\) (a state \(q \in Q \cup \{ \bot \} \) with \(\bot \) meaning halted, an input symbol \(a\), and work-tape symbols \(b_0, \dots , b_{k-1}\), all in \(\{ 0, 1, \sqcup \} \)), and let \(\tau \in \{ 0, \dots , k-1\} \). The symbol left under work head \(\tau \) is \(b_\tau \) if \(q = \bot \); if \(q \in Q\), it is the symbol in \(\{ 0, 1, \sqcup \} \) that the transition \(\delta (q, a, b)\) writes on work tape \(\tau \), or \(b_\tau \) when that transition writes nothing on tape \(\tau \).
Let \(\mu \) be a probability measure on a measurable space \(\Omega \), let \(\omega \mapsto A(\omega )\) be a family of \(k \times d\) real matrices indexed by \(\Omega \), and let \(\varepsilon \in \mathbb {R}\). Then the probability that the zero vector \(0 \in \mathbb {R}^d\) triggers the bad distortion event, namely that
satisfies
Let \(D\) be a digraph on a finite vertex set, and let \(\hat{D}\) denote its condensation: the digraph whose vertices are the strongly connected components (dicomponents) of \(D\), with an arc from one component to another precisely when \(D\) has an arc whose tail lies in the first and whose head lies in the second. Then \(\hat{D}\) contains no directed cycle; that is, for every dicomponent \(c\) and every closed walk in \(\hat{D}\) from \(c\) to \(c\), that walk is not a directed cycle — it does not have positive length together with distinct vertices apart from its coinciding endpoints.
Let \(G\) be a multigraph and let \(e \in E(G)\) be one of its edges. Then the multigraph \(G/e\) obtained by contracting \(e\) satisfies
where both edge counts are taken with multiplicity; in particular, contracting an edge never increases the number of edges and eliminates at least the contracted edge itself.
Let \(G\) be a multigraph and let \(e\) be an edge of \(G\) that is not a self-loop, so that its two endpoints are distinct. Writing \(G/e\) for the multigraph obtained by contracting \(e\), its vertex count satisfies
where the subtraction is truncated at \(0\).
Let \(D\) be a digraph on a finite vertex set, and let \(\breve{D}\) be its converse, obtained by reversing the orientation of every arc. Then for each vertex \(v\) the indegree and outdegree interchange:
where \(d^{-}\) counts the arcs with head \(v\) and \(d^{+}\) the arcs with tail \(v\).
Let \(D\) be a digraph with vertices \(u\) and \(v\), and let \(\breve{D}\) denote its converse, the digraph obtained by reversing the orientation of every arc. Then \(v\) is reachable from \(u\) in \(\breve{D}\) if and only if \(u\) is reachable from \(v\) in \(D\).
Fix a prime \(p\), let \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) be equipped with the symplectic form \(\omega \big((x,z),(x',z')\big) = \sum _{i} (x_i z'_i - z_i x'_i)\), and let \(S \le V\) be a subspace with symplectic orthogonal complement \(S^{\perp _\omega } = \{ v \in V : \omega (v,s) = 0 \text{ for all } s \in S\} \). Let \(M \subseteq \{ 0,1,\dots ,n-1\} \) be a set of coordinates, write \(V_M\) for the subspace of vectors supported on \(M\), and set \(S_M = S \cap V_M\) and \(S^{\perp }_M = S^{\perp _\omega } \cap V_M\). If \(M\) is correctable for \(S\) — meaning every vector of \(S^{\perp _\omega }\) supported on \(M\) already lies in \(S\), i.e. \(S^{\perp }_M \le S\) — then \(g(S,M) = \dim _{\mathbb {F}_p}(S^{\perp }_M) - \dim _{\mathbb {F}_p}(S_M) = 0\).
Let \(m\) be a multiset of unordered pairs (edges) drawn from a type \(\alpha \), and fix an element \(v \in \alpha \). For each edge \(e\), let \(\{ e\} _\ast \) denote the finite set of its endpoints — a singleton when \(e\) is a loop and a two-element set otherwise — and form the multiset obtained by concatenating \(\{ e\} _\ast \) over all \(e\) in \(m\), counting multiplicities of \(m\). Then the multiplicity of \(v\) in this concatenated multiset equals the number of edges \(e\) of \(m\) for which \(v\) is an endpoint of \(e\).
For a finite set \(S\) and an unordered pair \(e\), the predicate \(\texttt{Crosses}\; S\; e\) holds when some element of \(e\) lies in \(S\) and some element of \(e\) lies outside \(S\). This is a membership-based formulation independent of a choice of representative.
Let \(S\) be a finite subset of a type \(\alpha \), and let \(e\) be an unordered pair of elements of \(\alpha \); write \((x,y)\) for a chosen ordered representative of \(e\). Then \(e\) crosses \(S\) — that is, some element of \(e\) lies in \(S\) and some element of \(e\) lies outside \(S\) — if and only if either \(x \in S\) and \(y \notin S\), or \(x \notin S\) and \(y \in S\).
Let \(G\) be a multigraph and let \(C\) be a cut of \(G\) determined by a vertex subset \(S\). Then the multiset of edges crossing \(C\) equals the sub-multiset of edges \(e\) of \(G\) for which some endpoint of \(e\) lies in \(S\) while some endpoint of \(e\) lies outside \(S\), counted with multiplicity.
Fix \(N\) experts and a horizon \(T\), and let \(\ell \) be a loss sequence assigning to each round \(t\) and expert \(i\) a real loss \(\ell _t(i)\). For every expert \(i\), the cumulative loss through the first \(T\) rounds equals the total loss over all rounds: \(L_T(i) = \sum _{t} \ell _t(i)\), where \(t\) ranges over all \(T\) rounds.
Let \(\ell \) be a loss sequence for \(N\) experts over \(T\) rounds, and write \(L_t(i) = \sum _{s \lt t} \ell _s(i)\) for the cumulative loss of expert \(i\) through the first \(t\) rounds. Then for every round \(t \lt T\) and every expert \(i\),
Let \(G\) be a multigraph with vertex set \(V(G)\) and edge multiset \(E(G)\), and let \(C\) be a cut of \(G\), given by a nonempty proper subset \(S \subseteq V(G)\) with nonempty complement. Suppose \(e \in E(G)\) is an edge that does not cross \(C\), i.e. both endpoints of \(e\) lie on the same side of the partition. Then the cut induced on the contracted multigraph \(G/e\) by merging the endpoints of \(e\) has the same size as \(C\): the number of edges crossing the induced cut in \(G/e\), counted with multiplicity, equals the number of edges crossing \(C\) in \(G\).
Let \(D\) be a digraph with finite vertex set \(V\) and finite arc set \(A\), each arc \(a\) assigned a tail and a head in \(V\), and let \(F\) be a field. Write \(\mathcal{B} \le (A \to F)\) for the bond space of \(D\), the space of potential differences \(\delta p\) arising from functions \(p\) on \(V\). Suppose \(B\) is a basis matrix of \(\mathcal{B}\), so that the rows of \(B\) form a basis of \(\mathcal{B}\), and let \(S \subseteq A\). Then the columns of the restricted matrix \(B \mid S\), indexed by the arcs in \(S\), are linearly independent over \(F\) if and only if \(S\) is acyclic, i.e. contains no arc-cycle of \(D\).
Let \(V\) and \(A\) be finite sets of vertices and arcs, let each arc \(a\in A\) have a tail \(\mathrm{tail}(a)\in V\) and a head \(\mathrm{head}(a)\in V\), and let \(F\) be a field, so that the cycle space \(\mathcal{C}\le (A\to F)\) is defined. Let \(C\) be a basis matrix of \(\mathcal{C}\) — a matrix whose rows form a basis of \(\mathcal{C}\) — and let \(S\subseteq A\) be a finite set of arcs. Then the columns of the restriction \(C\mid S\) (the submatrix of \(C\) keeping exactly the columns labelled by elements of \(S\)) are linearly independent over \(F\) if and only if \(S\) contains no bond, that is, there is no bond \(B\) with \(B\subseteq S\).
g_B: the potential difference of a bond, p = indicator of S (B&M give it, p. 221).
Let B = [S, \(\bar{S}\)] be a bond of D. We define g_B by
g_B(a) = 1 if a \(\in \) (S, \(\bar{S}\)) = -1 if a \(\in \) (\(\bar{S}\), S) = 0 if a \(\notin \) B
It can be verified that g_B = \(\delta \)p where p(v) = 1 if v \(\in \) S and 0 if v \(\in \) \(\bar{S}\).
The bond space \(\mathcal{B}\): potential differences, i.e. g = M\(^{T}\) *\(_{v}\) p, B&M’s (12.2).
Given a function p on the vertex set V of D, we define the function \(\delta \)p on the arc set A by the rule that, if an arc a has tail x and head y, then
\(\delta \)p(a) = p(x) - p(y) (12.2)
If D is thought of as an electrical network with potential p(v) at v, then, by (12.2), \(\delta \)p represents the potential difference along the wires of the network. For this reason a function g on A is called a potential difference in D if g = \(\delta \)p for some function p on V. … the set \(\mathcal{B}\) of all potential differences in D is a vector space. … We shall see that each potential difference is a linear combination of potential differences associated with bonds. For this reason we refer to \(\mathcal{B}\) as the bond space of D.
Let \(D\) be a digraph with finite vertex set \(V\) and finite arc set \(A\), each arc \(a\) having a tail and a head in \(V\), and let \(F\) be a field. Let \(M\) be the oriented incidence matrix of \(D\) over \(F\), the \(V \times A\) matrix whose row \(m_v\) records, for each arc, whether \(v\) is its tail (\(+1\)), its head (\(-1\)), or neither (\(0\)). Then the bond space \(\mathcal{B} \subseteq F^{A}\) of \(D\) — the space of potential differences — coincides with the row space of \(M\), that is, with the subspace of \(F^{A}\) spanned by the rows \(\{ m_v : v \in V\} \).
f_C: the circulation of an oriented cycle (\(\pm \)1 on C\(^{+}\) / C\(^{-}\), 0 off C).
We associate with C the function f_C defined by
f_C(a) = 1 if a \(\in \) C\(^{+}\) = -1 if a \(\in \) C \ C\(^{+}\) = 0 if a \(\notin \) C
Clearly, f_C satisfies (12.1) and hence is a circulation.
The cycle space \(\mathcal{C}\): circulations, i.e. M *\(_{v}\) f = 0, B&M’s (12.1).
Let D be a digraph. A real-valued function f on A is called a circulation in D if it satisfies the conservation condition at each vertex:
f\(^{-}\)(v) = f\(^{+}\)(v) for all v \(\in \) V (12.1)
If we think of D as an electrical network, then such a function f represents a circulation of currents in D. … Thus the set of all circulations in D is a vector space. We denote this space by \(\mathcal{C}\). … We shall see later on that each circulation is a linear combination of the circulations associated with cycles. For this reason we refer to \(\mathcal{C}\) as the cycle space of D.
Let \(D\) be a directed graph with finite vertex set \(V\) and finite arc set \(A\), and let \(F\) be a field. Equip the space \(F^A\) of \(F\)-valued functions on the arcs with the standard dot-product bilinear form \(\langle f, g\rangle = \sum _{a \in A} f(a)\, g(a)\). Then the cycle space \(\mathcal{C} \subseteq F^A\) (the circulations) is exactly the orthogonal complement of the bond space \(\mathcal{B} \subseteq F^A\) (the potential differences) with respect to this form; that is, \(\mathcal{C} = \{ \, f \in F^A : \langle f, g\rangle = 0 \text{ for all } g \in \mathcal{B} \, \} \).
Let \(K\) be a field and let \(\iota \) be a finite index set, and equip the space of functions \(\iota \to K\) with the standard dot-product bilinear form. Then for all functions \(x, y \colon \iota \to K\), the value of this form at the pair \((x, y)\) is
Let \(D\) be a digraph with finite vertex set \(V\) and finite arc set \(A\), and let \(F\) be a field. If \(f \colon A \to F\) is a circulation on \(D\) — that is, \(f\) lies in the cycle space \(\mathcal{C}\) — and \(f\) is not identically zero, then there exists an arc-cycle \(c\) in \(D\) (a nonempty closed arc-sequence) all of whose arcs lie in the support of \(f\), i.e. every arc appearing in \(c\) satisfies \(f(a) \neq 0\).
Let \(D\) be a directed graph with finite vertex set \(V\) and finite arc set \(A\), and let \(F\) be a field. Suppose \(g \colon A \to F\) is a potential difference in \(D\) — that is, \(g\) lies in the bond space \(\mathcal{B}\) — and that \(g\) is not identically zero. Then there is a bond \(B\) of \(D\) (a minimal nonempty edge cut) all of whose arcs lie in the support of \(g\), i.e. \(B \subseteq \{ a \in A : g(a) \neq 0\} \).
Let \(D\) be a digraph with finite vertex set \(V\) and finite arc set \(A\), specified by its tail and head maps \(A \to V\), and let \(G\) be the underlying simple graph of \(D\). Working over the real field, the bond space \(\mathcal{B}\) of \(D\) — the space of potential differences on the arcs — satisfies
where \(\omega \) is the number of connected components of \(G\).
Let \(D\) be a digraph with finite vertex set \(V\) and finite arc set \(A\), given by its tail and head maps, and let \(G\) be its underlying simple graph. Then, working over \(\mathbb {R}\), the dimension of the cycle space \(\mathcal{C}\) of \(D\) is
where \(\omega \) is the number of connected components of \(G\).
Let \(D\) be a finite digraph with vertex set \(V\) and arc set \(A\), specified by its tail and head maps \(A \to V\), and let \(F\) be a field whose characteristic is a prime \(p\). Writing \(\mathcal{B}\) for the bond space and \(\mathcal{C}\) for the cycle space of \(D\) over \(F\), the intersection \(\mathcal{B} \cap \mathcal{C}\) has positive dimension over \(F\) if and only if \(p\) divides the number \(\tau (D)\) of spanning trees of \(D\).
The vertex set S of the fundamental bond of a \(\in \) T (\(\bar{T}\) + a contains a unique bond, B&M’s out-of-chapter Thm 2.6).
Analogously, if a is an arc of T, then \(\bar{T}\) + a contains a unique bond (see theorem 2.6). Let B_a denote this bond and g_a the potential difference corresponding to B_a, defined so that g_a(a) = 1.
The fundamental cycle of a \(\notin \) T (T + a contains a unique cycle). path_unique‘.
Let T be a maximal forest of D. Associated with T is a special basis matrix of \(\mathcal{C}\). If a is an arc of \(\bar{T}\), then T + a contains a unique cycle. Let C_a denote this cycle and let f_a denote the circulation corresponding to C_a, defined so that f_a(a) = 1.
Let \(G\) be a finite simple graph, and let \(M\) be its vertex–edge incidence matrix over \(\mathbb {Z}\), whose entry in row \(v\) and column \(e\) is \(1\) when \(v\) is an endpoint of the edge \(e\) and \(0\) otherwise. Then \(M\) is totally unimodular — that is, every square submatrix of \(M\) has determinant \(-1\), \(0\), or \(1\) — if and only if \(G\) is bipartite.
A basis matrix of a submodule W \(\le \) (A \(\to \) F): its rows are a basis of W.
A matrix B is called a basis matrix of \(\mathcal{B}\) if the rows of B form a basis for \(\mathcal{B}\); a basis matrix of \(\mathcal{C}\) is similarly defined.
Let \(D\) be a finite digraph with vertex set \(V\) and arc set \(A\), specified by its tail and head maps \(A \to V\), and let \(F\) be a field. Let \(T\) be a maximal forest of \(D\), and for each arc \(a \in T\) let \(g_a\) denote the potential difference of the fundamental bond of \(a\), that is, \(g_a = \delta p_a\) where \(p_a\) is the indicator function of the vertex set of that bond. Then the matrix over \(F\) whose rows are the functions \(g_a\), indexed by the arcs \(a \in T\), is a basis matrix of the bond space \(\mathcal{B}\) of \(D\); equivalently, the family \((g_a)_{a \in T}\) is a basis of \(\mathcal{B}\).
Let \(D\) be a digraph with finite vertex set \(V\) and finite arc set \(A\), and let \(F\) be a field. Fix a maximal forest \(T\) of \(D\). For each arc \(a \notin T\), the subgraph \(T + a\) contains a unique cycle, and we write \(f_a\) for the \(F\)-valued circulation associated with this fundamental cycle. Then the family \((f_a)_{a \notin T}\), viewed as the rows of a matrix whose row index ranges over the arcs outside \(T\), is a basis matrix of the cycle space \(\mathcal{C}\) of \(D\) over \(F\): these circulations are linearly independent and their span is all of \(\mathcal{C}\).
B is the tree-T basis matrix of the bond space (rows indexed by T).
The (\(\nu \)-\(\omega \)) \(\times \) \(\varepsilon \) matrix B whose rows are g_a, a \(\in \) T, is a basis matrix of \(\mathcal{B}\), called the basis matrix of \(\mathcal{B}\) corresponding to T.
C is the tree-T basis matrix of the cycle space (rows indexed by \(\bar{T}\)).
The (\(\varepsilon \)-\(\nu \)+\(\omega \)) \(\times \) \(\varepsilon \) matrix C whose rows are f_a, a \(\in \) \(\bar{T}\), is a basis matrix of \(\mathcal{C}\). This follows from the fact that each row is a circulation and that rank C = \(\varepsilon \) - \(\nu \) + \(\omega \) (because C | \(\bar{T}\) is an identity matrix). We refer to C as the basis matrix of \(\mathcal{C}\) corresponding to T.
Let a finite digraph be given by its tail and head maps on a finite arc set \(A\), and let \(T\) be a spanning tree of this digraph. If \(B\) is the integer basis matrix of the bond space corresponding to \(T\) — its rows indexed by the arcs of \(T\) — then \(B\) is unimodular; that is, every full square submatrix of \(B\) (a submatrix whose columns are picked out by an injective column selection) has determinant \(0\), \(+1\), or \(-1\).
Let \(T\) be a spanning tree of a finite digraph on vertex set \(V\) with arc set \(A\), and let \(C\) be the integer basis matrix of the cycle space corresponding to \(T\) — the matrix whose rows are indexed by the arcs outside \(T\) and whose columns are indexed by all of \(A\). Then \(C\) is unimodular: every full square submatrix of \(C\) formed by selecting a set of columns equal in number to the rows has determinant \(0\), \(+1\), or \(-1\).
Let \(D\) be a finite digraph with vertex set \(V\) and arc set \(A\), where each arc \(a\) has a tail and a head in \(V\), and let \(M\) be its oriented incidence matrix over \(\mathbb {Z}\): the rows are indexed by \(V\), the columns by \(A\), and the \((v,a)\) entry equals \(1\) if \(v\) is the tail of \(a\), equals \(-1\) if \(v\) is the head of \(a\), and equals \(0\) otherwise. Fix a vertex \(y \in V\), and let \(K\) be the matrix obtained from \(M\) by deleting the row indexed by \(y\). Then \(K\) is unimodular: for every injective map \(g\) from an index set into the columns of \(K\), the square submatrix of \(K\) formed by selecting the columns \(g\) (against all remaining rows) has determinant \(0\), \(1\), or \(-1\).
Let \(G\) be a simple graph on a finite vertex set \(V\), and let \((\mathrm{tail}, \mathrm{head})\) be an orientation of \(G\) on a finite arc set \(A\), so that each edge of \(G\) is directed from its tail to its head. Write \(M\) for the oriented incidence matrix of this orientation over \(\mathbb {Z}\), the \(V \times A\) integer matrix whose entry in row \(v\) and column \(a\) is \(1\) if \(v\) is the tail of \(a\), \(-1\) if \(v\) is the head of \(a\), and \(0\) otherwise. Then the graph Laplacian \(L(G)\) of \(G\), taken with integer entries, satisfies
The matrix-tree theorem (Kirchhoff, 1847). Let \(D\) be a connected digraph with vertex set \(V\) and arc set \(A\), and let \(M\) be its oriented incidence matrix, the \(V \times A\) matrix whose entry in row \(v\) and column \(a\) is \(1\) if \(v\) is the tail of \(a\), is \(-1\) if \(v\) is the head of \(a\), and is \(0\) otherwise. Fix any vertex \(y \in V\), and let \(K\) be the reduced incidence matrix obtained from \(M\) by deleting the row indexed by \(y\). Then the number \(\tau (D)\) of spanning trees of \(D\) satisfies
the equality holding in \(\mathbb {Z}\).
B&M’s oriented incidence matrix M (p. 222).
With each vertex v of D we associate the function m_v on A defined by
m_v(a) = 1 if a is a link and v is the tail of a = -1 if a is a link and v is the head of a = 0 otherwise
The incidence matrix of D is the matrix M whose rows are the functions m_v.
M | S — B&M’s column restriction (Matrix.submatrix).
If R is a matrix whose columns are labelled with the elements of A, and if S \(\subseteq \) A, we shall denote by R | S the submatrix of R consisting of those columns labelled with elements in S. If R has a single row, our notation is the same as the usual notation for the restriction of a function to a subset of its domain.
Let \(G\) be a finite digraph with vertex set \(V\) and arc set \(A\), and let \(T\) be a spanning tree of \(G\). Let \(B\) be the basis matrix of the bond space \(\mathcal{B}\) corresponding to \(T\), whose rows are indexed by the arcs of \(T\), and let \(C\) be the basis matrix of the cycle space \(\mathcal{C}\) corresponding to \(T\), whose rows are indexed by the arcs not in \(T\); both have integer entries and columns indexed by all of \(A\). Form the square integer matrix \(M\) with rows indexed by \(A\) by assigning to each arc of \(T\) its row of \(B\) and to each remaining arc its row of \(C\). Then the number \(\tau (G)\) of spanning trees of \(G\) satisfies \(\tau (G) = \det M\) or \(\tau (G) = -\det M\); that is, \(\tau (G) = \left\lvert \det M\right\rvert \).
Let \(G\) be a finite simple graph with vertex set \(V\), let \((\mathrm{tail}, \mathrm{head})\) be an orientation of \(G\), and fix a vertex \(y \in V\). Write \(L(G)\) for the integer Laplacian matrix of \(G\) (the vertex degree matrix minus the adjacency matrix), and let \(L(G)_{\widehat{y}}\) denote the reduced Laplacian obtained by deleting the row and the column indexed by \(y\), that is, the submatrix of \(L(G)\) whose rows and columns range over the vertices \(v \neq y\). Then the number of spanning trees \(\tau \) equals the determinant of this reduced Laplacian,
as an identity of integers, for every choice of the deleted vertex \(y\).
Let \(D\) be a finite digraph with vertex set \(V\) and arc set \(A\), specified by its tail and head maps \(A \to V\), and let \(T \subseteq A\) be a spanning tree of \(D\). Let \(B\) be the integer basis matrix of the bond space corresponding to \(T\), with rows indexed by the arcs of \(T\) and columns indexed by all of \(A\). Then the number \(\tau (D)\) of spanning trees of \(D\) satisfies
Let \(G\) be a directed graph with finite vertex set \(V\) and finite arc set \(A\), each arc assigned a tail and a head in \(V\), and let \(\tau \) denote the number of spanning trees of \(G\). Fix a spanning tree \(T \subseteq A\), and let \(C\) be the integer matrix over the cycle space of \(G\) that is the basis matrix corresponding to \(T\), so that its rows are indexed by the arcs \(a \notin T\) and its columns by the arcs of \(A\). Then
Let \(G\) be a finite simple graph on a vertex set \(V\), and let \((\mathrm{tail}, \mathrm{head})\) be an orientation of \(G\) with arc set \(A\), so that \(G\) is the underlying graph of the digraph \((\mathrm{tail}, \mathrm{head})\). Write \(L\) for the graph Laplacian of \(G\) over \(\mathbb {Z}\), the \(V \times V\) matrix whose diagonal entry at \(v\) is the degree of \(v\) and whose off-diagonal entry at \((u,v)\) is \(-1\) when \(u\) and \(v\) are adjacent and \(0\) otherwise. Fix vertices \(i, j \in V\) and a bijection \(e\) between the set of vertices other than \(i\) and the set of vertices other than \(j\), and let \(M\) be the submatrix of \(L\) whose rows are indexed by the vertices \(u \neq i\), whose columns are indexed by the vertices \(u \neq i\) as well, and whose \((u, u')\) entry is \(L_{u,\, e(u')}\). Then the number \(\tau \) of spanning trees of the digraph \((\mathrm{tail}, \mathrm{head})\) satisfies either \(\tau = \det M\) or \(\tau = -\det M\).
The de Bruijn digraph D_n: vertices are (n-1)-bit strings, arcs are left-shifts.
We define a digraph D_n as follows: the vertices of D_n are the (n-1)-digit binary numbers p\(_{1}\) p\(_{2}\) … p_{n-1} with p\(_{i}\) = 0 or 1. There is an arc with tail p\(_{1}\) p\(_{2}\) … p_{n-1} and head q\(_{1}\) q\(_{2}\) … q_{n-1} if and only if p_{i+1} = q\(_{i}\) for 1 \(\le \) i \(\le \) n-2; in other words, all arcs are of the form (p\(_{1}\) p\(_{2}\) … p_{n-1}, p\(_{2}\) p\(_{3}\) … p_n). In addition, each arc (p\(_{1}\) p\(_{2}\) … p_{n-1}, p\(_{2}\) p\(_{3}\) … p_n) of D_n is assigned the label p\(_{1}\) p\(_{2}\) … p_n.
Fix an integer \(n \ge 2\), and let \(D_n\) be the de Bruijn digraph whose vertices are the binary strings of length \(n-1\) and whose arcs run from each string \(p_1 p_2 \cdots p_{n-1}\) to every string of the form \(p_2 p_3 \cdots p_{n-1} b\) with \(b \in \{ 0,1\} \). Form the underlying undirected simple graph on the same vertex set, in which two distinct strings are adjacent precisely when an arc of \(D_n\) joins one to the other in either direction. Then this graph is connected: its vertex set is nonempty and any two vertices are linked by a path.
Let \(f\colon \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function on \(n\) variables, let \(k \le n\), and let \(acc\) be a partial assignment of the variables. Then the full decision tree for \(f\) that queries the variables \(k, k+1, \dots , n-1\) in order starting from \(acc\) has depth at most \(n - k\).
Let \(f \colon (\mathbb {F}_2)^n \to \{ \mathtt{true},\mathtt{false}\} \) be a Boolean function on \(n\) variables, let \(k \le n\), and let \(acc, x \colon \{ 0,\dots ,n-1\} \to \{ \mathtt{true},\mathtt{false}\} \) be two assignments that agree on every coordinate below \(k\), that is \(acc_i = x_i\) for all \(i \lt k\). Then the full decision tree for \(f\) that branches on variables \(k, k+1, \dots , n-1\) starting from the accumulator \(acc\) evaluates on input \(x\) to \(f(x)\).
Let \(S\) be a finite subset of \(\{ 0,1,\dots ,n-1\} \) and let \(i\) be one of these indices. Then the symmetric difference \(S \triangle \{ i\} \) satisfies
where \(\left\lvert S\right\rvert - 1\) denotes truncated subtraction on the natural numbers (so it is \(0\) when \(S\) is empty).
Let \(S\subseteq [n]\), let \(i\in [n]\), and let \(x\in \{ 0,1\} ^n\). Then the Walsh–Fourier character indexed by the symmetric difference \(S\mathbin {\triangle }\{ i\} \) satisfies
where \(\sigma \) is the sign encoding \(\sigma (x_i)=(-1)^{x_i}\).
The coefficient function \(\mathrm{coeffs}\, T : \mathcal{P}(\mathrm{Fin}\, n) \to \mathbb {R}\) defined by structural recursion on \(T\). A leaf labelled \(b\) assigns \(\mathrm{boolToSign}(b)\) to the empty frequency and \(0\) to all others; a branch on variable \(i\) with subtrees \(\mathrm{lo}, \mathrm{hi}\) assigns to \(S\) the value
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(S \subseteq \{ 0,1,\dots ,n-1\} \) be a set of variables whose cardinality \(\left\lvert S\right\rvert \) exceeds the depth of \(T\). Then the coefficient assigned to the frequency \(S\) vanishes, \(\mathrm{coeffs}\, T\, S = 0\).
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(S \subseteq \{ 0,\dots ,n-1\} \) be a frequency set. Then the coefficient \(\mathrm{coeffs}\, T\, S\) is a dyadic rational with denominator \(2^{\mathrm{depth}(T)}\): there exists an integer \(m \in \mathbb {Z}\) such that
Equivalently, every such coefficient is an integer multiple of \(2^{-\mathrm{depth}(T)}\).
Let \(T\) be a decision tree on \(n\) Boolean variables, and write \(\mathrm{coeffs}(T) : \mathcal{P}(\{ 0,\dots ,n-1\} ) \to \mathbb {R}\) for its recursively defined coefficient function. If \(k\) is a natural number with \(\mathrm{depth}(T) \le k\), then for every set \(S\) of variables the scaled coefficient \(\mathrm{coeffs}(T)(S)\cdot 2^{k}\) is an integer; that is, there exists \(m \in \mathbb {Z}\) with \(\mathrm{coeffs}(T)(S)\cdot 2^{k} = m\).
Extracts a deepest root-to-leaf path from a decision tree, at each branch descending into the deeper subtree (ties broken toward the high subtree). It returns the list of \((\text{queried variable}, \text{branch direction})\) pairs along that path.
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(f\colon \{ 0,1\} ^n\to \mathbb {R}\) be the \(\pm 1\)-valued function it computes, so that \(f(x)=\sigma (T(x))\) where \(\sigma \) sends \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\). Then \(f\) has Fourier degree at most the depth of \(T\): for every \(S\subseteq \{ 1,\dots ,n\} \) with \(\hat f(S)\ne 0\), one has \(\left\lvert S\right\rvert \le \operatorname {depth}(T)\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function, and let \(D(f)\) denote its decision-tree depth, the least \(d\) for which some decision tree of depth at most \(d\) computes \(f\). Let \(g : \{ 0,1\} ^n \to \mathbb {R}\) be its sign encoding, \(g(x) = \sigma (f(x))\), where \(\sigma \) sends \(0\) to \(1\) and \(1\) to \(-1\). Then \(g\) has degree at most \(D(f)\): every set \(S\) with \(\hat g(S) \neq 0\) satisfies \(\left\lvert S\right\rvert \le D(f)\).
For every Boolean function \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\), there is a decision tree \(T\) on the \(n\) Boolean variables whose depth is at most the decision-tree depth of \(f\) and which computes \(f\), meaning \(T\) evaluates to \(f(x)\) on every input \(x\).
Let \(T\) be a decision tree on \(n\) Boolean variables, let \(f\colon \{ 0,1\} ^n\to \{ -1,1\} \) be the \(\pm 1\)-valued function it computes under the sign encoding \(\mathrm{false}\mapsto 1\), \(\mathrm{true}\mapsto -1\), and let \(d\) be the depth of \(T\). Then for every set of coordinates \(S\subseteq [n]\) there is an integer \(m\) such that the Fourier–Walsh coefficient of \(f\) at frequency \(S\) satisfies
That is, every Fourier coefficient of \(f\) is an integer multiple of \(2^{-d}\).
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(\sigma \! \circ \! T\) denote its sign-encoded function \(x \mapsto \sigma (T(x))\) on the Boolean hypercube \(\{ 0,1\} ^n\). Then for every frequency set \(S \subseteq [n]\), the Fourier–Walsh coefficient of this function at \(S\) agrees with the recursively defined tree coefficient:
Fix \(n\), let \(c\) assign a real number \(c(S)\) to each subset \(S \subseteq [n]\), and consider the Boolean function \(f(x) = \sum _{S \subseteq [n]} c(S)\, \chi _S(x)\), where \(\chi _S\) is the Walsh–Fourier character of \(S\). Then for every \(T \subseteq [n]\) the Fourier–Walsh coefficient of \(f\) at frequency \(T\) recovers the corresponding coefficient, namely \(\hat f(T) = c(T)\).
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(T'\) be the tree obtained from \(T\) by negating the Boolean value at every leaf while leaving its branching structure unchanged. Then \(T'\) has the same depth as \(T\).
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(x \colon \{ 0,\dots ,n-1\} \to \{ \mathtt{false},\mathtt{true}\} \) be an assignment of the variables. Then the tree obtained from \(T\) by negating every leaf value evaluates on \(x\) to the negation of the value of \(T\) on \(x\); that is, its output equals \(\lnot \, (T(x))\).
Let \(T\) be a decision tree on \(n\) Boolean variables, computing the \(\pm 1\)-valued sign-encoded function \(x \mapsto \sigma (T(x))\), where \(\sigma \) is the sign encoding sending \(\mathrm{false}\) to \(1\) and \(\mathrm{true}\) to \(-1\). For every point \(x \in \{ 0,1\} ^n\),
where the sum ranges over all subsets \(S\) of \([n]\), \(c_T(S)\) denotes the recursively defined coefficient of \(T\) at the frequency \(S\), and \(\chi _S\) is the Walsh character of \(S\).
Let \(T\) be a decision tree on \(n\) Boolean variables, with size (number of leaves) \(\mathrm{size}(T)\) and depth \(\mathrm{depth}(T)\). Then \(\mathrm{size}(T) \le 2^{\mathrm{depth}(T)}\); that is, a decision tree of depth \(k\) has at most \(2^{k}\) leaves.
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(g \colon \{ 0,1\} ^n \to \{ -1,1\} \) be the sign-encoded function it computes. Then the number of frequencies \(S \subseteq [n]\) at which the Fourier–Walsh coefficient of \(g\) is nonzero satisfies
where \(\mathrm{size}(T)\) is the number of leaves of \(T\) and \(\mathrm{depth}(T)\) is its maximum root-to-leaf path length.
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(f\colon \{ 0,1\} ^n\to \{ -1,1\} \) be its sign-encoded tree function. Then the number of subsets \(S\subseteq [n]\) whose Fourier–Walsh coefficient \(\hat f(S)\) is nonzero is at most \(4^{\, \mathrm{depth}(T)}\).
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(g : \{ 0,1\} ^n \to \{ -1,1\} \) be the sign-encoded function it computes, sending an input \(x\) to \(\sigma (T(x))\) (with \(\mathrm{false}\mapsto 1\) and \(\mathrm{true}\mapsto -1\)). Then the sum of the absolute values of its Fourier–Walsh coefficients is bounded by the size of \(T\), i.e. its number of leaves:
where the sum ranges over all subsets \(S\) of \(\{ 1,\dots ,n\} \).
Let \(T\) be a decision tree on \(n\) Boolean variables, and for each subset \(S \subseteq \{ 0,\dots ,n-1\} \) let \(\mathrm{coeffs}\, T\, (S)\) denote its recursively defined Fourier coefficient at frequency \(S\). Then the total absolute mass of these coefficients is bounded by the number of leaves of \(T\):
Fix \(n \in \mathbb {N}\), let \(i\) be an element of \(\{ 0, 1, \dots , n-1\} \), and let \(g\) assign a real number to each subset of \(\{ 0, 1, \dots , n-1\} \). Then summing \(g\) over all subsets is unaffected by first toggling the membership of \(i\); that is,
where both sums range over all subsets \(S \subseteq \{ 0, 1, \dots , n-1\} \) and \(S \mathbin {\triangle } \{ i\} \) denotes the symmetric difference of \(S\) with the singleton \(\{ i\} \).
Fix \(n \in \mathbb {N}\) and an index \(i \in \{ 0, 1, \dots , n-1\} \). The map on subsets of \(\{ 0, 1, \dots , n-1\} \) that sends each set \(S\) to its symmetric difference \(S \triangle \{ i\} \) with the singleton \(\{ i\} \) is an involution: applying it twice returns every set to itself, so \(\bigl(S \triangle \{ i\} \bigr) \triangle \{ i\} = S\) for all \(S\).
Fix naturals \(n\) and \(d\), and let \(F_0, F_1, \dots , F_n\) be real coefficients defining the function \(F(x) = \sum _{j=0}^{n} F_j\, K_j^{(n)}(x)\). Suppose that \(F_0 \gt 0\), that \(F_j \ge 0\) for every \(j\) with \(1 \le j \le n\), and that \(F(x) \le 0\) for every integer \(x\) with \(d \le x \le n\). Then the maximum size of a binary code of length \(n\) and minimum distance \(d\) satisfies
Let \(\delta \) be a real number with \(0 \le \delta \le \tfrac 12\), and define \(\tau = \tfrac 12 - \sqrt{\delta (1-\delta )}\). Then \(\tau \) again lies in \([0,\tfrac 12]\) and applying the same rule to it recovers \(\delta \), that is, \(\tfrac 12 - \sqrt{\tau (1-\tau )} = \delta \). Hence the map \(x \mapsto \tfrac 12 - \sqrt{x(1-x)}\) is an involution of the interval \([0,\tfrac 12]\).
Let \(D\) be a finite digraph whose vertex set has at least two vertices. Then \(D\) is diconnected — every vertex is reachable from every other along a directed path — if and only if \(D\) is \(1\)-arc-connected; that is, if and only if for every nonempty proper subset \(S\) of the vertex set there is at least one arc of \(D\) with tail in \(S\) and head outside \(S\).
The 0/1 adjacency matrix of D.
Let v\(_{1}\), v\(_{2}\), …, v_\(\nu \) be the vertices of a digraph D. The adjacency matrix of D is the \(\nu \) \(\times \) \(\nu \) matrix A = [a_ij] in which a_ij is the number of arcs of D with tail v\(_{i}\) and head v\(_{j}\).
\(\varepsilon \), the number of arcs.
Throughout this chapter, D will denote a digraph and G its underlying graph. This is a useful convention; it allows us, for example, to denote the vertex set of D by V (since V = V(G)), and the numbers of vertices and arcs in D by \(\nu \) and \(\varepsilon \), respectively.
The condensation \(\hat{D}\): dicomponents contracted.
Let D\(_{1}\), D\(_{2}\), …, D_m be the dicomponents of D. The condensation \(\hat{D}\) of D is a directed graph with m vertices w\(_{1}\), w\(_{2}\), …, w_m; there is an arc in \(\hat{D}\) with tail w\(_{i}\) and head w\(_{j}\) if and only if there is an arc in D with tail in D\(_{i}\) and head in D\(_{j}\).
D is diconnected: every vertex reaches every other.
Two vertices are diconnected in D if each is reachable from the other. As in the case of connection in graphs, diconnection is an equivalence relation on the vertex set of D. The subdigraphs D[V\(_{1}\)], D[V\(_{2}\)], …, D[V_m] induced by the resulting partition (V\(_{1}\), V\(_{2}\), …, V_m) of V(D) are called the dicomponents of D. A digraph D is diconnected if it has exactly one dicomponent.
The induced subdigraph on S (kept on the same carrier to avoid subtype juggling).
A digraph D’ is a subdigraph of D if V(D’) \(\subseteq \) V(D), A(D’) \(\subseteq \) A(D) and \(\psi \)_{D’} is the restriction of \(\psi \)_D to A(D’). The terminology and notation for subdigraphs is similar to that used for subgraphs.
A directed cycle: a positive-length closed walk whose vertices (bar the repeated endpoint) are distinct.
The book defines it only by analogy — directed paths, directed cycles and directed tours are similarly defined — so the unfolding here (a closed directed walk of positive length whose vertices, apart from the coinciding endpoints, are distinct) follows §1.4’s cycle.
A directed path: a Quiver.Path with no repeated vertex.
A directed walk in D is a finite non-null sequence W = (v\(_{0}\), a\(_{1}\), v\(_{1}\), …, a_k, v_k), whose terms are alternately vertices and arcs, such that, for i = 1, 2, …, k, the arc a\(_{i}\) has head v\(_{i}\) and tail v_{i-1}. … A directed trail is a directed walk that is a trail; directed paths, directed cycles and directed tours are similarly defined.
D is k-arc-connected: every nonempty proper cut has \(\ge \) k outgoing arcs.
A nontrivial digraph D is k-arc-connected if, for every nonempty proper subset S of V, |(S, \(\bar{S}\))| \(\ge \) k.
D is an orientation of G — the chapter’s keystone gap.
With each digraph D we can associate a graph G on the same vertex set; corresponding to each arc of D there is an edge of G with the same ends. This graph is the underlying graph of D. Conversely, given any graph G, we can obtain a digraph from G by specifying, for each link, an order on its ends. Such a digraph is called an orientation of G.
D is a tournament: an orientation of the complete graph.
An orientation of a complete graph is called a tournament. The tournaments on four vertices are shown in figure 10.4. Each can be regarded as indicating the results of games in a round-robin tournament between four players; for example, the first tournament in figure 10.4 shows that one player has won all three games and that the other three have each won one.
\(\delta \)\(^{+}\), the minimum outdegree.
The same sentence as for \(\delta \)\(^{-}\): We denote the minimum and maximum indegrees and outdegrees in D by \(\delta \)\(^{-}\)(D), \(\Delta \)\(^{-}\)(D), \(\delta \)\(^{+}\)(D) and \(\Delta \)\(^{+}\)(D), respectively.
Let \(p\) be a prime and \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\), equipped with the symplectic form \(\omega \big((x,z),(x',z')\big) = \sum _{i} (x_i z'_i - z_i x'_i)\). For a coordinate subset \(M \subseteq \{ 0,\dots ,n-1\} \), let \(V_M \le V\) be the support submodule of vectors supported on \(M\), and for a subspace \(S \le V\) write \(S_M = S \cap V_M\). If \(S\) is isotropic, that is \(\omega (u,v) = 0\) for all \(u,v \in S\), then
where \(M^c = \{ 0,\dots ,n-1\} \setminus M\) is the complementary coordinate set.
Fix a prime \(p\) and an integer \(n \ge 0\), and let \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\). For a set \(E\) of coordinates, let \(r_E \colon V \to V\) be the restriction map that leaves each coordinate in \(E\) unchanged and sets every coordinate outside \(E\) to zero, and for a set \(C\) of coordinates let \(V_C \le V\) be the subspace of vectors supported on \(C\). Then for every subspace \(S \le V\),
where \(E^c\) denotes the complement of \(E\) among the \(n\) coordinates.
Fix a prime \(p\), an integer \(n\), and work in the symplectic space \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) equipped with the symplectic form \(\omega \big((x,z),(x',z')\big) = \sum _{i} (x_i z'_i - z_i x'_i)\). Let \(M\) be a subset of the coordinate set \(\{ 0,1,\dots ,n-1\} \), let \(V_M \le V\) be the support submodule of vectors whose coordinates vanish outside \(M\), and let \(r_M : V \to V_M\) be the restriction map that zeroes every coordinate outside \(M\). Then for any \(\mathbb {F}_p\)-subspace \(S \le V\), the symplectic orthogonal complement \(S^{\perp _\omega }\) satisfies
Fix a prime \(p\) and an integer \(n\), and let \(S\) be a subspace of the ambient space \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\). For a subset \(M \subseteq \{ 1,\dots ,n\} \) with complement \(M^c\), write \(S_M = S \cap V_M\) and \(S_{M^c} = S \cap V_{M^c}\) for the intersections of \(S\) with the corresponding support subspaces. Then
Fix a prime \(p\) and an integer \(n\), and let \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) be the space of pairs of coordinate vectors over the prime field \(\mathbb {F}_p\). For a subset \(C \subseteq \{ 1, \dots , n\} \), let \(V_C \subseteq V\) be the support submodule consisting of those pairs \((u, w)\) for which \(u_i = 0\) and \(w_i = 0\) at every index \(i \notin C\). Then \(V_C\) has dimension \(\dim _{\mathbb {F}_p}(V_C) = 2\lvert C\rvert \) over \(\mathbb {F}_p\).
Let \(p\) be prime and let \(S\) be an \(\mathbb {F}_p\)-subspace of the symplectic space \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\), with code distance \(d(S)\) and symplectic orthogonal complement \(S^{\perp _\omega }\). If \(E \subseteq \{ 0,\dots ,n-1\} \) is a set of erased coordinates with \(|E| \lt d(S)\), then \(E\) is correctable for \(S\): every \(v \in V_E \cap S^{\perp _\omega }\) already lies in \(S\).
Let \(\varphi \) be a DNF formula on \(n\) variables — a disjunction of terms, each term a conjunction of literals — and let \(x \in \{ 0,1\} ^n\) be an assignment. Form its De Morgan dual CNF by negating every literal of every term of \(\varphi \) and reading each resulting term as a clause (a disjunction of literals), so that the whole is a conjunction of clauses. Then this dual CNF evaluates on \(x\) to the negation of the value of \(\varphi \) on \(x\).
Let \(\varphi \) be a DNF formula on \(n\) variables, that is, a list of terms, each term a list of literals, and let its De Morgan dual CNF be obtained by negating every literal of every term. Then the dual CNF has the same width as \(\varphi \), where the width of a formula is the maximum number of literals appearing in any of its terms (equal to \(0\) when the formula is empty).
The decision-tree depth of \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) is the least \(d\) such that some decision tree of depth at most \(d\) computes \(f\); it is well defined since the full tree computes \(f\) at depth \(n\).
Let \(f\) be a Boolean function on \(n\) variables, and let \(\bar f\) denote its pointwise negation, so that \(\bar f(x) = 1 - f(x)\) for every input \(x\). Then \(f\) and \(\bar f\) have the same decision-tree depth: \(\operatorname {depth}_{\mathrm{DT}}(\bar f) = \operatorname {depth}_{\mathrm{DT}}(f)\).
Fix a natural number \(n\), and for a subset \(E \subseteq \mathrm{Fin}\, n\) write \(E^c = \mathrm{Fin}\, n \setminus E\) for its complement. Then for every subset \(M \subseteq \mathrm{Fin}\, n\), taking the complement twice recovers the original set: \((M^c)^c = M\).
Let \(G_0\) be a multigraph, and let \(C_0\) be a minimum cut of \(G_0\), so that the number of edges crossing \(C_0\) (counted with multiplicity), its size, is no larger than that of any other cut. Then for every natural number \(i\) with \(i \lt \left\lvert V(G_0)\right\rvert - 1\),
where \(\left\lvert V(G_0)\right\rvert \) is the vertex count and \(\left\lvert E(G_0)\right\rvert \) is the edge count of \(G_0\).
Let \(G\) be a loopless multigraph with finite vertex set \(V\) and edge multiset \(E\), and let \(k\) be a natural number. If every vertex \(v \in V\) has degree at least \(k\) in \(G\), then
where \(\left\lvert V\right\rvert \) is the number of vertices and \(\left\lvert E\right\rvert \) is the number of edges counted with multiplicity.
Edge connectivity \(\kappa \)’(G).
We … define the edge connectivity \(\kappa \)’(G) of G to be the minimum k for which G has a k-edge cut. If G is trivial, \(\kappa \)’(G) is defined to be zero. … G is said to be k-edge-connected if \(\kappa \)’(G) \(\ge \) k.
Given \(T \gt 0\) rounds of play in which the row player uses mixed strategy \(p_t\) and the column player uses \(q_t\), the empirical joint distribution assigns to each profile \((i,j)\) the time-averaged probability \(\frac{1}{T}\sum _{t=1}^{T} (p_t)_i \cdot (q_t)_j\).
Let \(G\) be a finite two-player game with column utility \(u_c\) over \(M\) row actions and \(N\) column actions, and consider \(T \gt 0\) rounds of play in which the row player uses the mixed strategy \(p_t\) and the column player the mixed strategy \(q_t\) at round \(t\). Fix \(\varepsilon \in \mathbb {R}\) and a pure column action \(j'\), and suppose the column player’s cumulative external regret against always playing \(j'\) is at most \(\varepsilon T\):
Then in the empirical joint distribution \(\sigma \) of the play, deviating to \(j'\) gains at most \(\varepsilon \) over the expected utility: \(\mathrm{colDevUtil}(\sigma , G, j') \le \mathrm{colExpUtil}(\sigma , G) + \varepsilon \). This is the column half of empiricalJoint_isApproxCCE, mirroring empiricalJoint_row_deviation_le.
Let \(G\) be a finite two-player game with \(M\) row actions and \(N\) column actions and column utility \(u_c(i,j)\), and fix a number of rounds \(T \ge 1\). Suppose that in each round \(t \in \{ 1,\dots ,T\} \) the row player plays a mixed strategy \(p_t\) over the \(M\) row actions and the column player plays a mixed strategy \(q_t\) over the \(N\) column actions, and let \(\sigma \) be the empirical joint distribution that assigns mass \(\frac{1}{T}\sum _{t=1}^{T} (p_t)_i\, (q_t)_j\) to each action profile \((i,j)\). Then, for every column action \(j'\), the column player’s expected utility from deviating to the pure action \(j'\) against the row marginal \(\sigma ^{\mathrm{row}}_i = \sum _{j} \sigma _{ij}\) is the time average of the round-by-round deviation utilities:
Consider a finite two-player game with \(M\) row actions and \(N\) column actions, whose column payoff is \(u_c(i,j)\) for the action profile \((i,j)\). Fix \(T \ge 1\) rounds in which the row player plays the mixed strategy \(p_t\) and the column player plays the mixed strategy \(q_t\), and let \(\sigma \) be the empirical joint distribution, which assigns to each profile \((i,j)\) the time-averaged probability \(\frac{1}{T}\sum _{t=1}^{T} (p_t)_i\, (q_t)_j\). Then the column player’s expected utility under \(\sigma \) equals the time-average of the per-round bilinear column payoffs:
Consider \(T \gt 0\) rounds of play between a row player with \(M\) actions and a column player with \(N\) actions, where in round \(t\) the row player uses the mixed strategy \(p_t\) over the row actions and the column player uses the mixed strategy \(q_t\) over the column actions. Form the empirical joint distribution \(\sigma \) assigning to each profile \((i,j)\) the time-averaged probability \(\frac{1}{T}\sum _{t=1}^{T} (p_t)_i (q_t)_j\). Then for each column action \(j\), the column marginal \(\sigma ^{\mathrm{col}}_j = \sum _i \sigma _{ij}\) equals the time-average of the column player’s weights on \(j\),
Let \(G\) be a finite two-player game with row utility \(u_r\) and column utility \(u_c\) over \(M\) row actions and \(N\) column actions, and consider \(T \gt 0\) rounds of play in which the row player uses the mixed strategy \(p_t\) and the column player uses the mixed strategy \(q_t\) at round \(t\). Fix \(\varepsilon \in \mathbb {R}\), and suppose that for every pure row deviation \(i'\),
and for every pure column deviation \(j'\),
Then the empirical joint distribution of the play, which assigns to each profile \((i,j)\) the probability \(\frac{1}{T}\sum _{t=1}^{T}(p_t)_i (q_t)_j\), is an \(\varepsilon \)-coarse correlated equilibrium of \(G\).
Let \(G\) be a finite two-player game with row utility \(u_r\) over \(M\) row actions and \(N\) column actions, and consider \(T \gt 0\) rounds of play in which the row player uses the mixed strategy \(p_t\) and the column player the mixed strategy \(q_t\) at round \(t\). Fix \(\varepsilon \in \mathbb {R}\) and a pure row action \(i'\), and suppose the row player’s cumulative external regret against always playing \(i'\) is at most \(\varepsilon T\):
Then in the empirical joint distribution \(\sigma \) of the play, deviating to \(i'\) gains at most \(\varepsilon \) over the expected utility: \(\mathrm{rowDevUtil}(\sigma , G, i') \le \mathrm{rowExpUtil}(\sigma , G) + \varepsilon \). This is the row half of empiricalJoint_isApproxCCE.
Let \(G\) be a finite two-player game with \(M\) row actions and \(N\) column actions, and let \(u_r(i,j)\) denote the row player’s utility. Fix an integer \(T \ge 1\), and suppose that in each round \(t \in \{ 1,\dots ,T\} \) the row player plays the mixed strategy \(p_t\) over the \(M\) row actions and the column player plays the mixed strategy \(q_t\) over the \(N\) column actions, giving rise to the empirical joint distribution \(\sigma \) with \(\sigma _{ij} = \tfrac {1}{T}\sum _{t=1}^{T} (p_t)_i\, (q_t)_j\). Then for every pure row action \(i'\), the row player’s deviation utility under \(\sigma \) equals the time average of the expected utilities of playing \(i'\) against each round’s column strategy:
Fix a finite two-player game with \(M\) row actions and \(N\) column actions, and write \(u_r(i,j)\) for the row player’s payoff at the pure profile \((i,j)\). Suppose that over \(T \gt 0\) rounds the row player plays a mixed strategy \(p_t\) and the column player a mixed strategy \(q_t\) in each round \(t\), and let \(\sigma \) be the empirical joint distribution given by \(\sigma _{ij} = \frac{1}{T}\sum _{t=1}^{T}(p_t)_i\, (q_t)_j\). Then the row player’s expected utility under \(\sigma \) equals the time-average of the per-round bilinear payoffs:
Fix \(T \ge 1\) rounds of play in which the row player uses mixed strategies \(p_1,\dots ,p_T\) over \(M\) actions and the column player uses mixed strategies \(q_1,\dots ,q_T\) over \(N\) actions, and let \(\sigma \) be the empirical joint distribution, which assigns to each profile \((i,j)\) the time-averaged probability \(\frac{1}{T}\sum _{t=1}^{T} (p_t)_i (q_t)_j\). Then for every row action \(i\), the row marginal \(\sigma ^{\mathrm{row}}_i = \sum _{j} \sigma _{ij}\) equals the time-average of the row player’s weight on \(i\):
Fix \(T \gt 0\) rounds of play in which, in round \(t\), the row player uses a mixed strategy \(p_t\) over \(\mathrm{Fin}\, M\) and the column player a mixed strategy \(q_t\) over \(\mathrm{Fin}\, N\); write \((p_t)_i\) and \((q_t)_j\) for the weights placed on actions \(i\) and \(j\). Let \(\sigma \) be the empirical joint distribution, whose value at each profile \((i,j)\) is the time average \(\sigma _{ij} = \frac{1}{T}\sum _{t=1}^{T} (p_t)_i\, (q_t)_j\). Then for every function \(f : \mathrm{Fin}\, M \times \mathrm{Fin}\, N \to \mathbb {R}\),
Given \(T \gt 0\) pure actions \(a_0, \dots , a_{T-1} \in \mathrm{Fin}\, n\), the empirical strategy is the mixed strategy assigning to each \(i\) the fraction of rounds on which action \(i\) was played:
Let \(G\) be a multigraph on a vertex type \(\alpha \), with vertex set \(V(G)\) and edge multiset \(E(G)\). If a point \(x \in \alpha \) occurs as an endpoint of some edge of \(G\) — that is, \(x\) lies in one of the unordered pairs making up \(E(G)\) — then \(x \in V(G)\).
Fix \(0 \lt \tau \lt 1/2\), and let \((t_n)\) be a sequence of natural numbers and \((a_n)\) a sequence of reals such that \(t_n/n \to \tau \) and, for all sufficiently large \(n\), the point \(a_n\) lies strictly below every real zero of the Krawtchouk polynomial \(K_{t_n}^{(n)}\). Then the truncated Christoffel–Darboux kernel evaluated at \((a_n, 0)\) grows at rate at most \(H(\tau )\), in the sense that
where \(H\) is the binary entropy function in base \(2\).
An \(\texttt{Episode}\ N\ T\) packages a learner policy, an adaptive adversary, and a finite loss sequence of length \(T\) for \(N\) experts, together with a consistency proof that the stored loss sequence equals the one generated by running \(\texttt{episodeLossSeq}\ T\ \mathrm{learner}\ \mathrm{adversary}\).
\(\texttt{Episode.generated}\ \mathrm{learner}\ \mathrm{adversary}\) is the canonical \(\texttt{Episode}\ N\ T\) constructed by taking the loss sequence to be exactly \(\texttt{episodeLossSeq}\ T\ \mathrm{learner}\ \mathrm{adversary}\); the consistency condition is satisfied by reflexivity.
Fix a number of experts \(N \ge 1\) and a number of rounds \(T\), and consider any episode for \(N\) experts over \(T\) rounds — that is, a learner policy, an adaptive adversary, and a loss sequence \(\ell \) of length \(T\) that equals the sequence generated by running that learner against that adversary. Then \(\ell \) is a valid loss sequence: for every round \(t \lt T\) and every expert \(i\), the generated loss satisfies \(0 \le \ell _t(i) \le 1\).
\(\texttt{episodeLossNat}\ \mathrm{learner}\ \mathrm{adversary}\ t : \mathrm{Fin}\, N \to \mathbb {R}\) is the loss vector produced at round \(t\) when the learner and adversary interact adaptively: the adversary sees the prefix of losses from rounds \(0, \ldots , t-1\) and the learner’s mixed strategy for round \(t\), then returns the loss vector for round \(t\). The definition is well-founded because the recursion only refers to strictly earlier rounds.
\(\texttt{episodeLossSeq}\ T\ \mathrm{learner}\ \mathrm{adversary} : \texttt{LossSeq}\ N\ T\) is the finite loss sequence realized over \(T\) rounds by running the learner against the adaptive adversary, obtained by restricting \(\texttt{episodeLossNat}\) to the first \(T\) rounds.
Fix a number of experts \(N \ge 1\) and a horizon \(T\), and let a learner policy play against an adaptive adversary over \(N\) experts. Then the episode loss sequence they generate over the first \(T\) rounds is valid: for every round \(t \lt T\) and every expert \(i\), the realized loss satisfies \(0 \le \ell _t(i) \le 1\).
Let \(e\) be an unordered pair of elements of a type \(\alpha \) with decidable equality, and let \((a, b)\) be the ordered pair chosen as a representative of \(e\). If an element \(z\) belongs to \(e\), then \(z = a\) or \(z = b\).
Let \(m, n, N\) be natural numbers with \(mn \lt N\), and let \(g_0, g_1, \dots , g_{N-1}\) be a sequence of \(N\) distinct integers. Then the sequence has an increasing subsequence of length \(m\) — indices \(i_1 \lt i_2 \lt \dots \lt i_m\) with \(g_{i_1} \lt g_{i_2} \lt \dots \lt g_{i_m}\) — or a decreasing subsequence of length \(n\) — indices \(j_1 \lt j_2 \lt \dots \lt j_n\) with \(g_{j_1} \gt g_{j_2} \gt \dots \gt g_{j_n}\).
Let \(C \le \mathcal{H}_n\) be a code on \(n\) qubits that is non-degenerate for \(t\) errors, meaning it satisfies the Knill–Laflamme condition \(P_C E^\dagger F P_C = \lambda _{EF} P_C\) for all Pauli strings \(E,F\) of weight at most \(t\) and moreover \(P_C E^\dagger F P_C = 0\) whenever \(E \neq F\). Then the error sphere \(\mathrm{ES}(C,t) = \bigvee _{\mathrm{wt}(p)\le t} \hat{p}(C)\) has dimension
where \(\mathcal{E}(n,t)\) is the set of Pauli strings of weight at most \(t\).
Let \(C \le \mathcal{H}_n\) be a code that is non-degenerate for \(t\)-error correction, and let \(E, F\) be Pauli strings on \(n\) qubits with weights \(\mathrm{wt}(E) \le t\) and \(\mathrm{wt}(F) \le t\) and \(E \neq F\). Then the images \(\hat{E}(C)\) and \(\hat{F}(C)\) of \(C\) under the corresponding Pauli operators are orthogonal subspaces of \(\mathcal{H}_n\); that is, \(\langle \hat{E}\, u,\, \hat{F}\, v\rangle = 0\) for all \(u, v \in C\).
Fix a real number \(p\) with \(0 \lt p \lt 1\) and an integer \(q \ge 2\), and let \(H_q\) denote the \(q\)-ary entropy function. Then there is a function \(\varepsilon \colon \mathbb {N}\to \mathbb {R}\) with \(\varepsilon (n) = o(n)\) such that, for all sufficiently large \(n\),
Let \(C\) be a code of length \(n\) over an alphabet \(\alpha \) — that is, a finite set of codewords, each a function \(\mathrm{Fin}\, n \to \alpha \) — and let \(d\) be a natural number that is the minimum distance of \(C\): there exist two distinct codewords of \(C\) at Hamming distance exactly \(d\), and every pair of distinct codewords of \(C\) is at Hamming distance at least \(d\). Then \(d \le n\).
Let \(\alpha \) be a finite nonempty field, let \(n\) and \(k\) be natural numbers with \(k \ge 1\), and let \(d\) be a positive integer. Regard an \(n \times k\) matrix \(G\) over \(\alpha \) as a map sending a message \(x \in \alpha ^{k}\) to the length-\(n\) codeword \(Gx\). Then the number of such matrices \(G\) for which some nonzero message \(x\) has \(\mathrm{wt}(Gx) \lt d\) is at most
where \(\mathrm{wt}\) denotes Hamming weight and \(B_{d-1}(\mathbf0)\) is the Hamming ball of radius \(d-1\) about the all-zero codeword of length \(n\).
Let \(\alpha \) be a finite field, and fix natural numbers \(n\), \(L\), \(M\) and a real number \(p\) with \(0 \le p \le 1\), \(1 \le L\), \(L \lt M\), and \(M \le \left\lvert \alpha \right\rvert ^n\). Suppose \(V\) is a natural number bounding the volume of every Hamming ball of radius \(\lfloor pn\rfloor \), in the sense that each such ball around an arbitrary word \(y \in \alpha ^n\) contains at most \(V\) codewords, and suppose that
Then there exists a code \(C\) of length \(n\) over \(\alpha \) consisting of exactly \(M\) codewords that is \((p,L)\)-list-decodable.
Let \(\alpha \) be a finite field, and let \(n\) and \(k\) be natural numbers. Fix a codeword \(v : \mathrm{Fin}\, n \to \alpha \) of length \(n\) and a codeword \(x : \mathrm{Fin}\, k \to \alpha \) of length \(k\). Then the set of matrices \(G \in \alpha ^{n \times k}\) satisfying \(Gx = v\) is finite.
Let \(p\) be a real number with \(0 \lt p\) and \(0 \lt 1 - p\), and set \(N_2 = \lceil 2/(p(1-p))\rceil + 1\). Then for every natural number \(n\) with \(N_2 \le n\), the floor \(\lfloor np\rfloor \) is strictly positive, i.e. \(\lfloor np\rfloor \ge 1\).
Let \(\alpha \) be a finite field with \(q = \left\lvert \alpha \right\rvert \) elements, and let \(n, k, d\) be natural numbers. Write \(B_{d-1}(\mathbf{0})\) for the Hamming ball of radius \(d-1\) about the all-zero codeword in \(\alpha ^n\), and suppose
Then there exists an \(n \times k\) generator matrix \(G\) over \(\alpha \) such that for every nonzero message \(x \in \alpha ^k\), the codeword \(Gx \in \alpha ^n\) has Hamming weight at least \(d\).
Let \(C\) be a code of length \(n\) over an alphabet \(\alpha \) whose minimum Hamming distance is \(d\), and suppose \(d \gt 0\). Write \(t = \left\lfloor \frac{d-1}{2} \right\rfloor \). Then for any two distinct codewords \(c_1, c_2 \in C\), the Hamming balls \(B_t(c_1)\) and \(B_t(c_2)\) of radius \(t\) centered at \(c_1\) and \(c_2\) are disjoint.
Let \(C\) be a code of length \(n\) over a finite alphabet \(\alpha \) whose minimum Hamming distance is \(d\), with \(d \gt 0\); that is, some two distinct codewords of \(C\) lie at Hamming distance exactly \(d\), and no two distinct codewords lie at distance less than \(d\). Then for any two distinct codewords \(c_1, c_2 \in C\), the Hamming balls of radius \(\lfloor (d-1)/2 \rfloor \) centered at \(c_1\) and \(c_2\) are disjoint: no codeword \(c'\) satisfies both \(d(c', c_1) \le \lfloor (d-1)/2 \rfloor \) and \(d(c', c_2) \le \lfloor (d-1)/2 \rfloor \).
Let \(\alpha \) be a finite alphabet, and regard a codeword of length \(n\) as a function \(\mathrm{Fin}\, n \to \alpha \), with the Hamming distance \(d(x,y)\) counting the coordinates in which two codewords differ. For every radius \(l\) and every center codeword \(c\), the Hamming ball \(B_l(c) = \{ c' : d(c',c) \le l\} \) has cardinality
where \(\left\lvert \alpha \right\rvert \) is the number of symbols in the alphabet.
Let \(\alpha \) be a finite alphabet with at least two symbols, write \(q = \left\lvert \alpha \right\rvert \) for its size, and let \(n\) be a natural number. Fix a real number \(p\) with \(0 \lt p \le 1 - 1/q\), and let
be the \(q\)-ary entropy function. Then for every codeword \(c \in \alpha ^n\), the Hamming ball of radius \(\lfloor np\rfloor \) about \(c\), namely \(B_{\lfloor np\rfloor }(c) = \{ c' \in \alpha ^n : d(c',c) \le \lfloor np\rfloor \} \), satisfies
Let \(\alpha \) be a finite alphabet with \(q = \left\lvert \alpha \right\rvert \gt 1\) symbols, and let \(C \subseteq \alpha ^n\) be a code (a finite set of length-\(n\) codewords) whose minimum Hamming distance is \(d\), with \(d \gt 0\). Then
Let \(\alpha \) be a finite field and let \(C\) be a linear code of length \(n\) over \(\alpha \), meaning there is a generator matrix \(G\) whose columns generate \(C\) (every \(G\)-image of a message word lies in \(C\), and every codeword of \(C\) is such an image). Suppose \(d\) is the minimum distance of \(C\), that is, some two distinct codewords of \(C\) lie at Hamming distance exactly \(d\) and no two distinct codewords lie closer. Then every nonzero codeword of \(C\) has Hamming weight at least \(d\), and some codeword of \(C\) has Hamming weight exactly \(d\), where the weight of a codeword is its Hamming distance from the all-zero codeword.
Let \(\alpha \) be a finite field, write \(q = \left\lvert \alpha \right\rvert \) for the number of its elements, and fix a real number \(p\) with \(0 \lt p \le 1 - \tfrac {1}{q}\) together with an integer \(L \ge 1\). In this regime one considers block codes over the alphabet \(\alpha \) that admit list decoding from a \(p\)-fraction of symbol errors with output lists of size at most \(L\), and the associated rate \(r\) they can achieve.
Fix an integer alphabet size \(q \ge 2\), a real number \(p\), and integers \(n\) and \(L \ge 1\). Set the rate
and define
Assume that \(0 \lt M \le q^n\) and \(L \lt M\). Then
Let \(\alpha \) be a finite field, let \(n\) and \(k \ge 1\) be natural numbers, and fix a nonzero codeword \(x \in \alpha ^{k}\) together with an integer \(d \gt 0\). Regard an \(n \times k\) matrix \(G\) over \(\alpha \) as chosen uniformly at random among all \(\left\lvert \alpha \right\rvert ^{nk}\) such matrices, so that the image \(Gx\) is a codeword of length \(n\). Then the probability that \(Gx\) has Hamming weight strictly less than \(d\) is at most the fraction of all codewords of length \(n\) that lie in the Hamming ball of radius \(d-1\) about the all-zero codeword:
For every natural number \(q\) and every real number \(p\),
where \(\log _q\) denotes the base-\(q\) logarithm. That is, the three terms of the \(q\)-ary entropy expression may be rewritten by commuting each product and absorbing the subtractions into the sign of the multiplier.
Let \(\alpha \) be a finite alphabet with \(q = \left\lvert \alpha \right\rvert \) symbols, and let \(p\) be a real number with \(0 \lt p \le 1 - 1/q\). Then the \(q\)-ary entropy of \(p\) is strictly positive:
Let \(\alpha \) be a finite alphabet with at least two symbols, and let \(C\) be a code of length \(n\) over \(\alpha \) — that is, a finite set of codewords \(c : \mathrm{Fin}\, n \to \alpha \). Suppose \(C\) has minimum distance \(d\), meaning that some two distinct codewords of \(C\) are at Hamming distance exactly \(d\) and no two distinct codewords are at distance less than \(d\). Then
Let \(a,b,n\in \mathbb {N}\) with \(a\gt 0\), \(b\gt 0\), and \(a+b=n\), and let \(c\gt 0\) be a real number. Suppose the factorial product \(a!\, b!\) admits the Stirling-type upper bound
Then the ratio \(n!/(a!\, b!)\) is bounded below by
Let \(\alpha \) be a finite field, and fix \(n \ge 0\) and \(k \ge 1\). Let \(x \in \alpha ^k\) be a nonzero vector. Then the distribution on \(\alpha ^n\) obtained by sampling an \(n \times k\) matrix \(G\) with entries in \(\alpha \) uniformly at random and returning the product \(Gx\) coincides with the uniform distribution on \(\alpha ^n\); that is, for every \(v \in \alpha ^n\),
Let \(q\) be an even integer with \(q \ge 2\), let \(k \ge 1\) and \(n \ge 2\) be integers, and let \(\delta , \varepsilon \gt 0\). Then there exists a constant \(C \gt 0\) such that
Let \(D\) be a finite digraph, let \(x \neq y\) be two of its vertices, and let \(l\) be a natural number. Suppose that \(d^{+}(x) - d^{-}(x) = l\) and \(d^{-}(y) - d^{+}(y) = l\), and that every other vertex \(v\) (that is, \(v \neq x\) and \(v \neq y\)) satisfies \(d^{+}(v) = d^{-}(v)\), where \(d^{+}\) and \(d^{-}\) denote outdegree and indegree. Then \(D\) contains \(l\) pairwise arc-disjoint directed paths from \(x\) to \(y\).
Let \(G\) be a finite simple graph. Then \(G\) admits an orientation \(D\) in which the in- and outdegrees at every vertex differ by at most one; that is, \(\left\lvert d^{+}(v) - d^{-}(v)\right\rvert \le 1\) for every vertex \(v\), where \(d^{+}(v)\) and \(d^{-}(v)\) denote the outdegree and indegree of \(v\) in \(D\).
Let \(D\) be a diconnected digraph on a finite vertex set, so that every vertex is reachable from every other along directed paths. Suppose the underlying simple graph of \(D\), obtained by forgetting the orientation of each arc, contains a cycle of odd length. Then \(D\) contains a directed cycle of odd length.
Let \(D\) be a strict digraph on a finite, nonempty vertex set, and let \(k\) be a positive integer such that \(\max \{ \delta ^{-}(D),\delta ^{+}(D)\} =k\), where \(\delta ^{-}(D)\) and \(\delta ^{+}(D)\) denote the minimum indegree and minimum outdegree of \(D\). Then \(D\) contains a directed cycle of length at least \(k+1\).
Let \(D\) be a digraph on a finite, nonempty vertex set \(V\) whose underlying undirected graph is connected. Then \(D\) admits a directed Euler tour—a closed directed trail that traverses every arc of \(D\) exactly once—if and only if every vertex \(v \in V\) is balanced, meaning its outdegree and indegree agree, \(d^{+}(v) = d^{-}(v)\).
Let \(D\) be a strict digraph — a loopless directed graph — on a finite, nonempty vertex set. Write \(\delta ^-(D)\) for the minimum indegree and \(\delta ^+(D)\) for the minimum outdegree, each taken over all vertices of \(D\). Then \(D\) contains a directed path — a directed walk in which no vertex is repeated — whose length (its number of arcs) is at least \(\max \{ \delta ^-(D),\, \delta ^+(D)\} \).
Let \(k \ge 1\) and \(d \ge 0\) be integers. Then there is a probability space \((\Omega , \mathcal{F}, \mu )\) carrying a measurable map \(\omega \mapsto A(\omega )\), valued in the \(k \times d\) real matrices, with the following properties: each entry \(A_{ij}\), for \(1 \le i \le k\) and \(1 \le j \le d\), has law \(\mathcal{N}(0, 1/k)\) (mean zero and variance \(1/k\)); for each row index \(i\) the family of entries \((A_{ij})_{j}\) is jointly independent; and the family of rows \(\bigl((A_{ij})_{j}\bigr)_{i}\) is jointly independent.
Let \(\mu \) be a probability measure on \(\Omega \), let \(A : \Omega \to \mathbb {R}^{k\times d}\) be a measurable random matrix, let \(\varepsilon \in \mathbb {R}\), and let \(V\) be a finite subset of \(\mathbb {R}^d\). Suppose that for every ordered pair \(u, v \in V\) the bad event \(\{ \omega \mid \texttt{BadSingle}\, \varepsilon \, (A\, \omega )\, (u-v)\} \) is measurable and has probability at most \(2\exp (-k\varepsilon ^2/8)\), and that the union-bound failure probability satisfies
Then there exists a linear map \(f : \mathbb {R}^d \to \mathbb {R}^k\) which is an \(\varepsilon \)-JL embedding of \(V\), i.e. \((1-\varepsilon )\| u-v\| ^2 \le \| f u - f v\| ^2 \le (1+\varepsilon )\| u-v\| ^2\) for all \(u, v \in V\).
Let \(G\) be a finite graph on a nonempty vertex set that admits an Eulerian trail — a walk traversing every edge of \(G\) exactly once — and let \(k\) be a natural number with \(2k \le \kappa '(G)\). Then \(G\) has an orientation that is \(k\)-arc-connected.
Let \(G\) be a finite simple graph with maximum degree \(\Delta (G)\). Then \(G\) admits an orientation \(D\) — a digraph on the same vertex set whose underlying graph is \(G\) and in which no two vertices are joined by arcs in both directions — such that every directed path in \(D\), that is, every directed walk that repeats no vertex, has length at most \(\Delta (G)\), where the length of a path is the number of arcs it traverses.
Let \(G\) be a simple graph on a finite vertex set that admits a proper colouring with \(k\) colours. Then \(G\) has an orientation \(D\) — a digraph whose underlying graph is \(G\) and in which no pair of vertices is joined by arcs in both directions — such that every directed path in \(D\) (a directed walk that repeats no vertex) has length at most \(k-1\), that is, passes through at most \(k\) vertices.
Let \(D\) be a digraph on a finite vertex set \(V\) with \(n = |V|\), and suppose \(D\) is acyclic: it contains no directed cycle, that is, no closed directed walk of positive length whose vertices, apart from the coinciding endpoints, are distinct. Then \(V\) admits a topological ordering: there is a bijection \(f \colon V \to \{ 0, 1, \dots , n-1\} \) such that \(f(u) \lt f(v)\) for every arc of \(D\) from \(u\) to \(v\).
For every \(\eta \gt 0\) and every \(x \in [0,1]\), the value \(e^{-\eta x}\) lies below the chord joining the endpoints of the graph of \(t \mapsto e^{-\eta t}\) over \([0,1]\); that is, the inequality
holds, the right-hand side being the affine function equal to \(1\) at \(x = 0\) and to \(e^{-\eta }\) at \(x = 1\).
Fix a prime \(p\) and a subset \(C \subseteq \{ 1,\dots ,n\} \) of coordinates, and let \(V_C \subseteq \mathbb {F}_p^n \times \mathbb {F}_p^n\) be the submodule of pairs of vectors whose support is contained in \(C\). Then extending back the coordinate functions of a supported vector recovers that vector exactly: for every \(x \in V_C\), applying the restriction map \(V_C \to (C \to \mathbb {F}_p) \times (C \to \mathbb {F}_p)\) and then the extension-by-zero map back into \(V_C\) returns \(x\). Equivalently, the extension map is a left inverse of the restriction map.
Fix natural numbers \(n\), \(t\), and \(d\) with \(t + 1 \le n\), and let \(a\) be a real number lying strictly below every real zero of the Krawtchouk polynomial \(K_t^{(n)}\). Suppose that the truncated Christoffel–Darboux kernel satisfies \(\Phi _{n,t}(a, x) \le 0\) for every integer \(x\) with \(d \le x \le n\). Then the maximum size of a binary code of length \(n\) and minimum distance \(d\) is bounded by
Let \(p\) be a prime and equip the \(2n\)-dimensional space \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) over \(\mathbb {F}_p\) with the symplectic form \(\omega \big((x,z),(x',z')\big) = \sum _{i=1}^{n} (x_i z'_i - z_i x'_i)\). If \(S \le V\) is an isotropic subspace, meaning \(\omega (u,v) = 0\) for all \(u, v \in S\), then \(\dim _{\mathbb {F}_p} S \le n\).
Let \(p\) be a prime and equip the \(2n\)-dimensional space \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) over the prime field \(\mathbb {F}_p\) with the symplectic form \(\omega \big((x,z),(x',z')\big) = \sum _{i=1}^{n}(x_i z'_i - z_i x'_i)\). Then for every subspace \(S \le V\), its symplectic orthogonal complement \(S^{\perp _\omega } = \{ v \in V \mid \omega (v,s) = 0 \text{ for all } s \in S\} \) has dimension
Let \(p\) be a prime and \(\mathbb {F}_p\) the field with \(p\) elements, and for a natural number \(n\) set \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\). Then \(V\) is a vector space over \(\mathbb {F}_p\) of dimension
Let \(S\) be an isotropic submodule of the symplectic space \(V=\mathbb {F}_p^n\times \mathbb {F}_p^n\), and let \(M\subseteq \{ 0,\dots ,n-1\} \) be a set of coordinates with complement \(M^c\). Writing \(S_M=S\cap V_M\) and \(S_{M^c}=S\cap V_{M^c}\) for the parts of \(S\) supported on \(M\) and on its complement, one has
Let \(S\) be an isotropic subspace of the symplectic space \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) (with \(p\) prime), and let \(M \subseteq \{ 0,1,\dots ,n-1\} \) be an erasure set that is correctable for \(S\). Writing \(M^c\) for the complement of \(M\), one has
Fix a prime \(p\) and equip \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) with the symplectic form \(\omega \big((x,z),(x',z')\big) = \sum _{i=0}^{n-1}\big(x_i z'_i - z_i x'_i\big)\). For a set of coordinates \(M \subseteq \{ 0,\dots ,n-1\} \), let \(V_M \le V\) be the subspace of vectors supported on \(M\), and for a subspace \(W \le V\) write \(W_M = W \cap V_M\). Let \(S \le V\) be isotropic, meaning \(\omega (u,v) = 0\) for all \(u,v \in S\), let \(S^{\perp }\) denote its symplectic orthogonal complement, and set \(g(S,M) = \dim _{\mathbb {F}_p}\! \big(S^{\perp }\cap V_M\big) - \dim _{\mathbb {F}_p}(S_M)\). Then, writing \(M^c\) for the complementary set of coordinates,
Let \(p\) be prime, and equip the symplectic space \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) with its standard alternating form \(\omega \). If \(S \le V\) is an isotropic subspace and \(M \subseteq \{ 0,1,\dots ,n-1\} \) is any set of coordinates, then
where \(M^c\) is the complement of \(M\) and \(S_{M^c} = S \cap V_{M^c}\) consists of the vectors of \(S\) supported on \(M^c\).
Let \(p\) be prime and let \(S\) be an \(\mathbb {F}_p\)-subspace of the symplectic space \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\). Suppose \(B\) and \(C\) are disjoint subsets of the coordinate set \(\{ 0,\dots ,n-1\} \), and that \(B\) is correctable for \(S\). Then
where \(g(S,M) = \dim _{\mathbb {F}_p}\! \bigl(S^{\perp }_M\bigr) - \dim _{\mathbb {F}_p}\! \bigl(S_M\bigr)\).
A finite two-player game with \(M\) row actions and \(N\) column actions, specified by a row utility function \(u_r : \mathrm{Fin}\, M \times \mathrm{Fin}\, N \to \mathbb {R}\) and a column utility function \(u_c : \mathrm{Fin}\, M \times \mathrm{Fin}\, N \to \mathbb {R}\). Both players are utility maximizers.
Let \(D\) be a loopless finite digraph on a vertex set of size \(n \gt 2\), and suppose that every vertex \(v\) satisfies \(d^{-}(v) \ge n/2\) and \(d^{+}(v) \ge n/2\), where \(d^{-}(v)\) and \(d^{+}(v)\) denote the indegree and outdegree of \(v\). Then \(D\) contains a directed cycle of length \(n\); that is, a closed directed walk of positive length whose vertices, apart from the coinciding endpoints, are distinct, and which passes through every vertex of \(D\).
Let \(\mathtt{eval} : \mathit{Hyp} \to X \to \mathsf{Bool}\) assign a Boolean label to each input under each hypothesis, let \(H\) be a finite class of hypotheses, and fix a target hypothesis \(\mathtt{target} \in H\) that supplies the correct label \(\mathtt{eval}(\mathtt{target}, x)\) for every input \(x\). Then, on any finite sequence \(xs\) of inputs, the number of mistakes made by the Halving Algorithm run from the initial version space \(H\) is at most \(\lfloor \log _2 \left\lvert H\right\rvert \rfloor \):
Let \(\mathtt{eval} : \mathit{Hyp} \to X \to \mathsf{Bool}\) be an evaluation map, let \(V \subseteq \mathit{Hyp}\) be a finite version space, let \(x \in X\) be an input, and let \(y \in \mathsf{Bool}\) be its true label. If the Halving prediction on \(x\) is wrong, that is \(\mathtt{predict}(\mathtt{eval}, V, x) \ne y\), then the updated version space contains at most half of the hypotheses:
\(\mathtt{mistakes}(\mathtt{eval}, \mathtt{target}, V, xs)\) is the total number of prediction errors made by the Halving Algorithm when it processes the input list \(xs\), using the version space \(V\), with the true label at each step supplied by \(\mathtt{target}\). It is defined recursively: the empty list yields zero mistakes, and for a list \(x :: xs\) the count is \(1\) if the algorithm mispredicts on \(x\) (else \(0\)), plus the mistakes on the remainder using the updated version space.
Let \(\mathtt{eval} : \mathit{Hyp} \to X \to \mathsf{Bool}\) be an evaluation map, let \(V\) be a finite version space of hypotheses, and fix a target hypothesis \(\mathtt{target} \in V\) whose predictions supply the true label at each step. Then for every input sequence \(xs\), the total number of mistakes made by the Halving Algorithm — which predicts by majority vote over the current version space and, after each input, retains only those hypotheses that agreed with the observed label — satisfies
The Halving Algorithm predicts on input \(x\) by taking a majority vote of the current version space \(V\): it returns \(\mathsf{true}\) whenever the number of hypotheses in \(V\) that predict \(\mathsf{true}\) is at least as large as the number that predict \(\mathsf{false}\) (ties are broken in favour of \(\mathsf{true}\)).
Let \(\mathtt{eval} : \mathit{Hyp} \to X \to \mathsf{Bool}\) be an evaluation map, let \(V \subseteq \mathit{Hyp}\) be a version space, fix an input \(x \in X\), and let \(\mathtt{target} \in V\) be a hypothesis. Then \(\mathtt{target}\) remains in the version space obtained by updating \(V\) on \(x\) with the label that \(\mathtt{target}\) itself assigns, that is, \(\mathtt{target} \in \mathtt{update}(\mathtt{eval}, V, x, \mathtt{eval}\, \mathtt{target}\, x)\).
After observing that the correct label for input \(x\) is \(y\), the version space is updated to \(\mathtt{update}(\mathtt{eval}, V, x, y) = \mathtt{voteFor}(\mathtt{eval}, V, x, y)\), i.e. only the hypotheses that correctly predict \(y\) on \(x\) are retained.
Fix an evaluation map \(\mathtt{eval} : \mathit{Hyp} \to X \to \mathsf{Bool}\), a finite version space \(V \subseteq \mathit{Hyp}\), an input \(x \in X\), and a label \(y \in \mathsf{Bool}\). Then the updated version space is contained in the original one: \(\mathtt{update}(\mathtt{eval}, V, x, y) \subseteq V\).
Given an evaluation map \(\mathtt{eval} : \mathit{Hyp} \to X \to \mathsf{Bool}\), a version space \(V \subseteq \mathit{Hyp}\), an input \(x \in X\), and a label \(y \in \mathsf{Bool}\), \(\mathtt{voteFor}(\mathtt{eval}, V, x, y)\) is the subset of \(V\) consisting of all hypotheses that predict label \(y\) on input \(x\).
Let \(\mathrm{eval}\) be a Boolean evaluation map assigning to each hypothesis and each input a truth value, let \(V\) be a finite version space of hypotheses, and let \(x\) be an input. Every hypothesis in \(V\) predicts either \(\mathsf{false}\) or \(\mathsf{true}\) on \(x\), so the number of hypotheses in \(V\) predicting \(\mathsf{false}\) on \(x\) and the number predicting \(\mathsf{true}\) on \(x\) sum to the total number of hypotheses in \(V\):
Fix an integer \(k \ge 1\), and let \(Y\) be a real-valued random variable on a probability space \((\Omega , \mu )\) whose law is the centered Gaussian \(\mathcal{N}(0, 1/k)\) of mean \(0\) and variance \(1/k\). Then the centered square \(Y^2 - 1/k\) satisfies the Bernstein MGF condition on \((\Omega , \mu )\) with quadratic-growth parameter \(c = 2/k^2\) and radius \(t_{\max } = k/4\).
Let \(X\) be a real random variable on a measure space \((\Omega ,\mu )\) that is sub-Gaussian with variance proxy \(\sigma \ge 0\), i.e. \(e^{tX}\) is \(\mu \)-integrable and \(\mathbb {E}_\mu [e^{tX}] \le \exp (\sigma t^2/2)\) for every \(t \in \mathbb {R}\). Then for every constant \(c \in \mathbb {R}\), the scaled variable \(\omega \mapsto c\, X(\omega )\) is sub-Gaussian with variance proxy \(c^2\sigma \):
Let \(\omega = (\omega _{ij})\) be a random \(k \times d\) real array whose entries are independent and identically distributed, each taking the values \(-1\) and \(+1\) with probability \(\tfrac {1}{2}\) apiece. Then for every index \(i \in \{ 1,\dots ,k\} \) and \(j \in \{ 1,\dots ,d\} \), the single entry \(\omega _{ij}\) is sub-Gaussian with variance proxy \(1\); that is, its moment generating function satisfies \(\mathbb {E}\! \left[e^{t\, \omega _{ij}}\right] \le e^{t^2/2}\) for every \(t \in \mathbb {R}\).
Let \(\mathrm{Rad}\) denote the Rademacher distribution on \(\mathbb {R}\), placing mass \(1/2\) at each of \(-1\) and \(+1\), and regard the identity map as the real random variable \(X\) with law \(\mathrm{Rad}\). Then \(X\) is sub-Gaussian with variance proxy \(1\): its moment generating function satisfies
Fix an integer \(k \ge 1\) and a dimension \(d\), let \(x \in \mathbb {R}^d\) (with its Euclidean norm), and fix a row index \(i \in \{ 1,\dots ,k\} \). Draw the entries of the Rademacher matrix \(A(\omega )_{ij} = \omega _{ij}/\sqrt{k}\) by making all \(kd\) signs \(\omega _{ij}\) independent Rademacher variables. Then the real random variable
has a sub-Gaussian moment generating function with variance proxy \(\lVert x\rVert ^2/k\): for every \(t \in \mathbb {R}\),
Fix natural numbers \(k \gt 0\) and \(d\), a vector \(x \in \mathbb {R}^d\), a row index \(i\), and a column index \(j\). Consider the joint Rademacher product measure on the space of \(k \times d\) real arrays \(\omega = (\omega _{ij})\), under which every entry is an independent \(\pm 1\) Rademacher variable (equal to \(-1\) or \(+1\) with probability \(\tfrac 12\) each). Then the scaled entry \(\omega \mapsto (x_j/\sqrt{k})\, \omega _{ij}\) has a sub-Gaussian moment generating function with variance proxy \(x_j^2/k\); that is, its moment generating function is finite and satisfies
Let \(H\) be a \(3\)-uniform hypergraph on the vertex set \(\mathrm{Fin}\, n\), so that every edge of \(H\) has exactly three vertices, and let \(d\) be a positive integer. Call a two-element set \(Q = \{ u,v\} \) of vertices *heavy* if its degree exceeds \(d\), that is, if more than \(d\) edges of \(H\) contain \(Q\), and let \(P\) denote the collection of all heavy pairs. Then
For a hypergraph \(H\) on \(\mathrm{Fin}\, n\) and a threshold \(d\), the set of heavy pairs
realised as the collection of two-element sets \(\{ u,v\} \) whose degree in \(H\) exceeds \(d\).
Let \(G\) be a finite two-player zero-sum game with \(M \ge 2\) row actions and \(N \ge 1\) column actions, with payoff matrix \((A_{ij})\) satisfying \(0 \le A_{ij} \le 1\); the row player maximises and the column player minimises. Fix an integer \(T \gt 0\) and a learning rate \(\eta \gt 0\). Then there exist a mixed row strategy \(p\) and a mixed column strategy \(q\) such that, for every pure row \(i \in \{ 1,\dots ,M\} \) and every pure column \(j \in \{ 1,\dots ,N\} \),
Fix \(N \gt 1\) experts and \(T\) rounds with \(T \gt 0\), and suppose \(2\ln N \le T\). Let \(\ell \) be a valid loss sequence for \(N\) experts over \(T\) rounds, so that every individual loss satisfies \(0 \le \ell _t(i) \le 1\). Running the Hedge algorithm with the optimal learning rate \(\eta ^* = \sqrt{2\ln N / T}\), the average regret against the best expert in hindsight is bounded by
Fix \(N \ge 1\) experts and \(T\) rounds, and let \(\ell \) be a valid loss sequence, so that \(0 \le \ell _t(i) \le 1\) for every round \(t\) and expert \(i\). Then for any learning rate \(\eta \) with \(0 \lt \eta \le 1\), the regret of the Hedge algorithm satisfies
Fix \(N \ge 1\) experts and a horizon of \(T\) rounds, and run the Hedge algorithm with a learning rate \(\eta \gt 0\). For every valid loss sequence \(\ell \) — one whose losses satisfy \(0 \le \ell _t(i) \le 1\) for all rounds \(t\) and experts \(i\) — the regret of Hedge is bounded by
Fix \(N \ge 1\) experts, a horizon of \(T\) rounds, and a learning rate \(\eta \gt 0\). Let the Hedge algorithm with rate \(\eta \) — the learner policy whose round-\(t\) distribution is the exponential-weights distribution \(p_t(i) \propto \exp (-\eta L_t(i))\) built from the past realized cumulative losses — play against an arbitrary adaptive adversary, whose loss vectors have every coordinate in \([0,1]\) and may depend on the full history of losses and on the learner’s current distribution. Let \(\ell \) be the loss sequence realized over the \(T\) rounds of this interaction. Then the regret of Hedge, its cumulative expected loss minus the least cumulative loss \(L^* = \min _{i} L_T(i)\) of any single expert in hindsight, satisfies
Fix \(N \ge 1\) experts, a horizon of \(T\) rounds, and a learning rate \(\eta \gt 0\). Consider a full interaction in which the learner runs the Hedge algorithm with learning rate \(\eta \) — playing at each round \(t\) the exponential-weights distribution \(p_t(i) \propto \exp (-\eta L_t(i))\) over the experts, where \(L_t(i)\) is the cumulative loss of expert \(i\) over rounds \(0,\dots ,t-1\) — against an adaptive adversary who, having observed the past loss vectors and the learner’s current distribution, returns a loss vector with every coordinate in \([0,1]\); let \(\ell \) be the loss sequence so generated. Writing \(\widehat{L}_T = \sum _{t=0}^{T-1}\sum _{i} p_t(i)\, \ell _t(i)\) for the learner’s expected cumulative loss and \(L^\ast = \min _{i} L_T(i)\) for the loss of the best expert in hindsight, the regret satisfies
Consider the Hedge algorithm run on \(N \gt 1\) experts over \(T \gt 0\) rounds against a valid loss sequence \(\ell \), that is, one with \(0 \le \ell _t(i) \le 1\) for every round \(t\) and expert \(i\). Suppose \(2\ln N \le T\), and run Hedge with the learning rate \(\eta ^* = \sqrt{2\ln N / T}\). Then the regret—the learner’s total expected loss minus the loss of the best single expert in hindsight—satisfies
Consider the Hedge algorithm run against a valid loss sequence for \(N \gt 1\) experts over \(T \ge 1\) rounds, so that every individual loss lies in \([0,1]\). If the learning rate is set to the tight optimal value \(\eta _{\mathrm{tight}} = \sqrt{8\ln N / T}\), then the regret against the best expert in hindsight satisfies
Fix a number of experts \(N \ge 1\), a horizon \(T\), a learning rate \(\eta \in \mathbb {R}\), a loss sequence \(\ell \) for \(N\) experts over \(T\) rounds, and a round \(t\). Then the Hedge distribution at round \(t\) sums to one over all experts,
so it is a probability distribution on the \(N\) experts.
Fix a learning rate \(\eta \in \mathbb {R}\), and let \(\ell \) be a valid loss sequence for \(N \ge 1\) experts over \(T\) rounds, so that \(0 \le \ell _t(i) \le 1\) for every round \(t\) and expert \(i\). Then at each round \(t\) the expected loss of the learner under the Hedge distribution satisfies \(\widehat{\ell }_t \le 1\).
Let \(\ell \) be a valid loss sequence for \(N \ge 1\) experts over \(T\) rounds, fix any learning rate \(\eta \in \mathbb {R}\), and let \(t\) be one of the rounds. Then the learner’s expected loss \(\widehat{\ell }_t = \sum _{i=1}^{N} p_t(i)\, \ell _t(i)\) under the Hedge distribution is nonnegative, \(\widehat{\ell }_t \ge 0\).
\(\texttt{hedgePolicy}\ N\ \eta \) is the \(\texttt{LearnerPolicy}\) that implements the Hedge algorithm with learning rate \(\eta \gt 0\): at each round \(t\) and for each loss history \(h\), the weights are given by \(\texttt{hedgeDist}\ \eta \ (\text{\texttt{LossHistory.toLossSeq}}\ h)\ t\), the exponential-weights distribution derived from the past realized losses.
At each round \(t \in \mathrm{Fin}\, T\), Hedge’s prediction in the original decision space is the weighted average
where \(w^{(\eta )}_{t,i} = \texttt{hedgeDist}\, \eta \, L\, t\, i\) are the Hedge weights computed from the induced loss table \(L\) up to (but not including) round \(t\), and \(p_t(i)\) is expert \(i\)’s prediction at round \(t\).
Fix \(N \ge 1\) experts, a horizon of \(T\) rounds, a learning rate \(\eta \in \mathbb {R}\), and a convex set \(S \subseteq \mathbb {R}\). Let \(\ell : \mathbb {R}\times \Omega \to \mathbb {R}\) be a loss function, let \(e_t(i) \in \mathbb {R}\) be expert \(i\)’s prediction at round \(t\), and let \(o_t \in \Omega \) be the outcome at round \(t\); write \(L_{t,i} = \ell (e_t(i), o_t)\) for the induced loss table and let \(q_t(i)\) denote the Hedge distribution over the experts computed from \(L\) at round \(t\). Suppose every prediction \(e_t(i)\) lies in \(S\), and that for each round \(t\) the map \(x \mapsto \ell (x, o_t)\) is convex on \(S\). Then at every round \(t\) the loss of Hedge’s weighted-average prediction \(\widehat{x}_t = \sum _{i} q_t(i)\, e_t(i)\) is at most its expected loss under \(q_t\):
Let \(S \subseteq \mathbb {R}\) be a convex set, fix any learning rate \(\eta \in \mathbb {R}\), a loss function \(\ell : \mathbb {R}\to \Omega \to \mathbb {R}\), expert predictions \(p_t(i)\), and a sequence of outcomes, and suppose that every expert prediction lies in \(S\), that is \(p_t(i) \in S\) for all rounds \(t\) and all experts \(i\). Then for every round \(t\), Hedge’s weighted-average prediction \(\widehat{x}_t = \sum _{i} w^{(\eta )}_{t,i}\, p_t(i)\) — formed from the Hedge distribution of the loss table induced by \(\ell \), the predictions, and the outcomes up to round \(t\) — again lies in \(S\).
Let \(S \subseteq \mathbb {R}\) be a convex set, let \(\eta \gt 0\) be a learning rate, and consider a prediction problem with \(N \ge 1\) experts over \(T\) rounds, an outcome \(o_t\) at each round \(t\), and a loss function \(\ell (x, o)\). Suppose that every expert prediction \(p_t(i)\) lies in \(S\), that for each round \(t\) the map \(x \mapsto \ell (x, o_t)\) is convex on \(S\), and that the induced loss sequence is valid, so that each expert’s per-round loss \(\ell (p_t(i), o_t)\) lies in \([0,1]\). Then the cumulative loss \(\mathcal{L}^{\mathrm{pred}}_T = \sum _{t} \ell (\widehat{x}_t, o_t)\) of Hedge’s weighted-average predictions \(\widehat{x}_t\) exceeds the best expert loss in hindsight \(L^* = \min _{i} \sum _{t} \ell (p_t(i), o_t)\) by at most
Consider prediction with expert advice using \(N \ge 2\) experts over \(T \ge 1\) rounds, where predictions take values in a convex set \(S \subseteq \mathbb {R}\). Let \(\ell : \mathbb {R}\times \Omega \to \mathbb {R}\) be a loss function, let \(p_t(i) \in S\) be the prediction of expert \(i\) at round \(t\), and let \(o_t \in \Omega \) be the outcome at round \(t\); assume that for every round \(t\) the map \(x \mapsto \ell (x, o_t)\) is convex on \(S\), and that all induced losses \(\ell (p_t(i), o_t)\) lie in \([0,1]\). Run the Hedge algorithm with learning rate \(\eta _{\mathrm{tight}} = \sqrt{8\ln N / T}\), predicting at each round the Hedge weighted average \(\widehat{x}_t = \sum _i p_t(i)\, w_{t}(i) \in S\), where the weights \(w_t\) form the Hedge distribution built from the induced losses through the first \(t\) rounds. Then the cumulative loss of these predictions exceeds the best expert’s loss in hindsight by at most \(\sqrt{\tfrac {T}{2}\ln N}\):
The actual cumulative prediction loss of Hedge is
where \(\widehat{x}_t\) is the Hedge weighted-average prediction at round \(t\). This differs from hedgeCumLoss: here the loss is applied to the averaged prediction, not to a weighted average of expert losses.
Fix a learning rate \(\eta \in \mathbb {R}\), a loss function \(\ell : \mathbb {R}\to \Omega \to \mathbb {R}\), expert predictions \(p : \mathrm{Fin}\, T \to \mathrm{Fin}\, N \to \mathbb {R}\) (with \(N \ge 1\)), and a sequence of outcomes \(o : \mathrm{Fin}\, T \to \Omega \). Suppose \(S \subseteq \mathbb {R}\) is convex, every expert prediction \(p_t(i)\) lies in \(S\), and for each round \(t\) the map \(x \mapsto \ell (x, o_t)\) is convex on \(S\). Then the cumulative loss of Hedge’s weighted-average predictions is at most the cumulative Hedge loss computed on the induced loss sequence:
where \(\widehat{x}_t = \sum _{i} p_t(i)\, p_t(i)\) is the Hedge weighted-average prediction at round \(t\) and \(p_t\) is the Hedge distribution at that round.
The sequence of column best responses generated online against Hedge: at time \(t\), given the history \(a_0, \dots , a_{t-1}\) of prior column choices, the column player responds with \(a_t = \mathrm{bestColumn}(G, p_t)\), where \(p_t\) is the prefix Hedge mixed strategy. This is defined by well-founded recursion on \(t\).
Let \(G\) be a finite zero-sum game with \(M \ge 1\) row actions and \(N \ge 1\) column actions, let \(\eta \in \mathbb {R}\) be a learning rate, and let \(T\) be a horizon. Write \(j_s = \texttt{hedgeResponseNat}(G, \eta , s)\) for the online column response at round \(s\), and let \(\ell = \texttt{ZeroSumGame.toLossSeq}(G, (j_s)_{s \lt T})\) be the loss sequence it induces for the row player. Then for every round \(t \lt T\) and every fixed column \(j\),
i.e. the played column \(j_t\) gives the row player, mixing according to the Hedge distribution at round \(t\), no more payoff than any other column.
Fix a learning rate \(\eta \in \mathbb {R}\), a loss sequence \(\ell \) for \(N\) experts over \(T\) rounds, a time \(t\), and an expert \(i\). Then the unnormalized Hedge weight \(w_t(i) = \exp \! \bigl(-\eta \cdot L_t(i)\bigr)\), where \(L_t(i) = \sum _{s \lt t} \ell _s(i)\) is the cumulative loss of expert \(i\) through the first \(t\) rounds, is strictly positive.
Fix a learning rate \(\eta \in \mathbb {R}\) and a loss sequence \(\ell \) for \(N\) experts over \(T\) rounds. For every round \(t \in \mathrm{Fin}\, T\) and every expert \(i \in \mathrm{Fin}\, N\), the unnormalized Hedge weights satisfy
Let \(L\) be a normal form \((3,\delta ,\varepsilon )\)-decodable code with message length \(k \ge 1\) and block length \(n \ge 1\), and suppose its total number of constraints is positive, \(m \ge 1\). If the advantage parameter satisfies \(\varepsilon \le 1/2\), then there exist a message \(b : \mathrm{Fin}\, k \to \mathbb {Z}\) and an assignment \(x : \mathrm{Fin}\, n \to \mathbb {Z}\), each taking values in \(\{ -1, 1\} \), such that the XOR polynomial of \(L\) satisfies
Let \(p_1,\dots ,p_n\) be a probability distribution, so \(p_i \ge 0\) for every \(i\) and \(\sum _{i=1}^n p_i = 1\), and let \(\ell _1,\dots ,\ell _n \in [0,1]\) be losses. Write \(L = \sum _{i=1}^n p_i\, \ell _i\) for the expected loss. Then for every \(\eta \gt 0\),
Let \(n \ge 1\) and let \(p = (p_1, \dots , p_n)\) be a probability vector, so that \(p_i \ge 0\) for every \(i\) and \(\sum _{i} p_i = 1\). Let \(\ell _1, \dots , \ell _n \in [0,1]\), and let \(\eta \gt 0\). Writing \(L = \sum _{i} p_i \ell _i\) for the average loss, one has
Let \(L\) be a normal form locally decodable code with message length \(k\) and block length \(n\), whose message indices \(i : \mathrm{Fin}\, k\) carry the \(3\)-uniform matchings \(H_i\) on \(\mathrm{Fin}\, n\), and let \(d \ge 1\) be a natural number. Then there is a family of hypergraphs \(H'_i\) on \(\mathrm{Fin}\, n\), one for each \(i\), such that every \(H'_i\) consists of edges of \(H_i\) (that is, \(H'_i \subseteq H_i\)) and is again \(3\)-uniform and a matching, while the union \(\bigcup _{i} H'_i\) satisfies the pair-degree bound \(d\): every pair of distinct vertices \(u \neq v\) is contained in at most \(d\) of its edges.
Given a loss function \(\ell : \mathbb {R} \to \Omega \to \mathbb {R}\), expert predictions \(p : \mathrm{Fin}\, T \to \mathrm{Fin}\, N \to \mathbb {R}\), and a sequence of outcomes \(o : \mathrm{Fin}\, T \to \Omega \), the induced loss sequence is the abstract loss table \(L : \mathrm{LossSeq}\, N\, T\) defined by \(L_{t,i} = \ell (p_t(i),\, o_t)\), i.e. expert \(i\) at round \(t\) incurs the loss its own prediction receives under the realized outcome.
Let \(\alpha \) be a linearly ordered type, let \(x \in \alpha \), and let \(xs\) be a finite list of elements of \(\alpha \). Then the list obtained by timed ordered insertion of \(x\) into \(xs\) is a permutation of \(x :: xs\), the list with \(x\) prepended to \(xs\).
Let \(\alpha \) be a linearly ordered type, let \(x \in \alpha \), and let \(\mathit{xs}\) be a sorted list over \(\alpha \). Then the list returned by the timed ordered insertion of \(x\) into \(\mathit{xs}\) — the return value of that computation, discarding its time cost — is again sorted.
Let \(\alpha \) be a linearly ordered type and let \(\mathit{xs}\) be a finite list of elements of \(\alpha \). Then the list returned by the timed insertion sort applied to \(\mathit{xs}\) is sorted, and it is a permutation of \(\mathit{xs}\).
Let \(v\) be a nonnegative real number with \(v \ne 0\), and let \(N(0,v)\) denote the centred Gaussian probability measure on \(\mathbb {R}\) with variance \(v\). If \(t \in \mathbb {R}\) satisfies \(2tv \lt 1\), then the function \(y \mapsto e^{t y^2}\) is integrable with respect to \(N(0,v)\).
Let \(\mathrm{Rad} = \tfrac {1}{2}\delta _{-1} + \tfrac {1}{2}\delta _{+1}\) be the Rademacher distribution on \(\mathbb {R}\). Then every function \(f \colon \mathbb {R}\to \mathbb {R}\) is integrable with respect to \(\mathrm{Rad}\).
Fix a point \(a \in \mathbb {R}\), and let \(\delta _a\) denote the Dirac measure concentrated at \(a\). Then every function \(f : \mathbb {R}\to \mathbb {R}\) is integrable with respect to the finite measure \(\tfrac {1}{2}\, \delta _a\).
Let \(\omega = (\omega _{ij})\) be a \(k \times d\) array whose entries are independent Rademacher variables, each taking the values \(-1\) and \(+1\) with probability \(\tfrac 12\) apiece. Then for every pair of indices \(i \in \{ 1,\dots ,k\} \) and \(j \in \{ 1,\dots ,d\} \), the entry \(\omega _{ij}\) has mean zero: \(\mathbb {E}[\omega _{ij}] = 0\).
Let \(v \gt 0\) and write \(N(0,v)\) for the centered Gaussian distribution on \(\mathbb {R}\) with mean \(0\) and variance \(v\). Then for every real number \(t\) with \(2tv \lt 1\),
Let \(Z\) be a measurable real random variable on a probability space \((\Omega , \mu )\) which is strictly sub-Gaussian with variance proxy \(\sigma ^2\): for every real \(s\), \(e^{sZ}\) is \(\mu \)-integrable and \(\mathbb {E}_\mu [e^{sZ}] \le e^{\sigma ^2 s^2/2}\). Then for every \(t \ge 0\) with \(2\sigma ^2 t \lt 1\), the function \(e^{tZ^2}\) is \(\mu \)-integrable and
the same bound that holds for \(Z \sim N(0, \sigma ^2)\).
Let \(Z\) be a standard Gaussian random variable, \(Z \sim N(0,1)\), and let \(s \in \mathbb {R}\) satisfy \(2s \lt 1\). Then
Let \(\mathrm{Rad} = \tfrac {1}{2}\delta _{-1} + \tfrac {1}{2}\delta _{+1}\) be the Rademacher distribution on \(\mathbb {R}\). Then the mean of the identity function against \(\mathrm{Rad}\) vanishes,
Let \(Z\) be a measurable real random variable on a probability space \((\Omega , \mu )\) which is strictly sub-Gaussian with variance proxy \(\sigma ^2 \gt 0\) (for every real \(s\), \(e^{sZ}\) is integrable and \(\mathbb {E}_\mu [e^{sZ}] \le e^{\sigma ^2 s^2/2}\)), and suppose \(\mathbb {E}_\mu [Z^2] = \sigma ^2\). Then \(Z^4\) is \(\mu \)-integrable and
(For a Gaussian the exact value is \(3\sigma ^4\).)
Fix integers \(k \gt 0\) and \(d \ge 0\), and let the entries \(\omega _{ij}\), for \(i \in \{ 1,\dots ,k\} \) and \(j \in \{ 1,\dots ,d\} \), be independent Rademacher random variables, each taking the values \(-1\) and \(+1\) with probability \(1/2\). Form the \(k \times d\) random matrix \(M\) with entries \(M_{ij} = \omega _{ij}/\sqrt{k}\). Then for every fixed vector \(x \in \mathbb {R}^d\) and every row index \(i \in \{ 1,\dots ,k\} \), the \(i\)-th coordinate of the image \(Mx\) has expectation zero:
Let \(\mathrm{Rad} = \tfrac {1}{2}\delta _{-1} + \tfrac {1}{2}\delta _{+1}\) be the Rademacher distribution on \(\mathbb {R}\), placing equal mass \(1/2\) at \(-1\) and at \(+1\). Then its second moment equals one:
Fix integers \(k \gt 0\) and \(d \ge 0\), a vector \(x \in \mathbb {R}^d\), and a row index \(i \in \{ 1, \dots , k\} \). Draw \(\omega \) from the joint Rademacher measure on \(\mathrm{Rad\Omega }(k,d)\), form the Rademacher matrix \(A(\omega ) \in \mathbb {R}^{k \times d}\), and let \((A(\omega )x)_i\) denote the \(i\)-th coordinate of the image of \(x\) under the linear map \(A(\omega )\). Then the second moment of this coordinate is
A joint distribution \(\sigma \) is an \(\varepsilon \)-coarse correlated equilibrium (\(\varepsilon \)-CCE) of game \(G\) if each player’s gain from any unilateral pure deviation is at most \(\varepsilon \): for all \(i'\), \(\mathrm{rowDevUtil}(\sigma , G, i') \le \mathrm{rowExpUtil}(\sigma , G) + \varepsilon \), and similarly for the column player.
A joint distribution \(\sigma \) is a coarse correlated equilibrium (CCE) of game \(G\) if neither player can improve their expected utility by committing in advance to any fixed pure action: for all \(i'\), \(\mathrm{rowDevUtil}(\sigma , G, i') \le \mathrm{rowExpUtil}(\sigma , G)\), and for all \(j'\), \(\mathrm{colDevUtil}(\sigma , G, j') \le \mathrm{colExpUtil}(\sigma , G)\).
Let \(G\) be a finite two-player game with \(M\) row actions and \(N\) column actions, and let \(\sigma \) be a joint distribution over the action profiles \((i,j)\) that is a coarse correlated equilibrium of \(G\); that is, neither player can increase their expected utility under \(\sigma \) by committing in advance to a fixed pure action, so that for every row action \(i'\) the row player’s utility from deviating to \(i'\) (against the column marginal of \(\sigma \)) is at most their expected utility under \(\sigma \), and likewise the column player’s utility from deviating to any fixed action \(j'\) (against the row marginal of \(\sigma \)) is at most their expected utility under \(\sigma \). Then for every real \(\varepsilon \ge 0\), the distribution \(\sigma \) is an \(\varepsilon \)-coarse correlated equilibrium of \(G\): each player’s gain from any such unilateral pure deviation exceeds zero by at most \(\varepsilon \).
A linear map \(f : \mathbb {R}^d \to _L[\mathbb {R}] \mathbb {R}^k\) is an \(\varepsilon \)-JL embedding of a finite set \(V\) if, for every ordered pair \((u, v) \in V \times V\), the pair \((f(u), f(v))\) satisfies JLDistortion \(\varepsilon \; u\; v\).
Let \(\alpha \) be a linearly ordered type and let \(\mathit{xs}\) be a finite list of elements of \(\alpha \). Then \(\mathit{xs}\) is sorted — in the sense that each element is at most the one following it — if and only if \(\mathit{xs}\) is a chain for the relation \(\le \), meaning that \(x \le y\) holds for every pair of consecutive entries \(x, y\) of \(\mathit{xs}\).
Let \(\mathit{xs}\) be a finite list over a linearly ordered set. Then \(\mathit{xs}\) is sorted in the adjacent-pairs sense if and only if it is sorted in the usual pairwise sense with respect to \(\le \); that is, if and only if \(x_i \le x_j\) for every pair of indices \(i \lt j\).
Let \(D\) be a digraph on a finite, nonempty vertex set \(V\). Then \(D\) is unilateral if and only if there exist vertices \(u, v \in V\) and a directed walk from \(u\) to \(v\) — a finite sequence of edges, each traversed in the direction it points — whose vertices include every vertex of \(V\). In other words, \(D\) is unilateral exactly when it admits a spanning directed walk.
Let \(k \ge 1\) and let \(A\) be a random \(k \times d\) matrix whose entries \(A_{ij}\) are independent Gaussian random variables, each distributed as \(\mathcal{N}(0, 1/k)\). Then for every fixed vector \(x \in \mathbb {R}^d\) and every \(\varepsilon \) with \(0 \lt \varepsilon \lt 1\), the probability that \(A\) distorts the squared norm of \(x\) by more than a factor \(\varepsilon \) satisfies
Let \(k \ge 1\) be an integer, let \(x \in \mathbb {R}^d\) be a nonzero vector, and fix \(\varepsilon \) with \(0 \lt \varepsilon \lt 1\). Let \(A\) be the \(k \times d\) matrix whose \(kd\) entries are independent, each taking the values \(+1/\sqrt{k}\) and \(-1/\sqrt{k}\) with probability \(1/2\) apiece. Then the squared norm \(\lVert Ax\rVert ^2\) concentrates about \(\lVert x\rVert ^2\), with the probability of a deviation exceeding \(\varepsilon \lVert x\rVert ^2\) bounded by
The bound is Achlioptas’s \(\pm 1\)-entry variant of the Johnson–Lindenstrauss lemma, with the Dasgupta–Gupta exponent.
Let \((\Omega ,\mu )\) be a probability space, let \(A\colon \Omega \to \mathbb {R}^{k\times d}\) be a random matrix with \(k\ge 1\) rows, and let \(x\in \mathbb {R}^d\) be a nonzero vector. Write \((Ax)_i\) for the \(i\)-th coordinate of the image \(Ax\in \mathbb {R}^k\), and suppose that the \(k\) coordinates \((Ax)_1,\dots ,(Ax)_k\) are measurable and mutually independent, that each is sub-Gaussian with variance proxy \(\lVert x\rVert ^2/k\) — meaning \(e^{t(Ax)_i}\) is \(\mu \)-integrable and \(\mathbb {E}_\mu \! \bigl[e^{t(Ax)_i}\bigr]\le \exp \! \bigl((\lVert x\rVert ^2/k)\, t^2/2\bigr)\) for every \(t\in \mathbb {R}\) — and that each has second moment \(\mathbb {E}_\mu \! \bigl[(Ax)_i^2\bigr]=\lVert x\rVert ^2/k\). Then for every \(\varepsilon \) with \(0\lt \varepsilon \lt 1\),
Fix an integer \(k \ge 1\), a dimension \(d\), a probability space \((\Omega , \mu )\), and a vector \(x \in \mathbb {R}^d\). Let \(A : \Omega \to \mathbb {R}^{k \times d}\) be a random matrix, and for each outcome \(\omega \) write \(y(\omega ) = A(\omega )\, x \in \mathbb {R}^k\) for the image of \(x\) under the random linear map, with coordinates \(y_1(\omega ), \dots , y_k(\omega )\). Suppose each coordinate \(y_i\) is measurable, the coordinates \(y_1, \dots , y_k\) are independent, and, for some constants \(c \gt 0\) and \(t_{\max }\), each centered squared coordinate \(y_i^2 - \lVert x\rVert ^2/k\) satisfies the Bernstein moment generating function condition with parameters \((c, t_{\max })\). Then for every \(s\) with \(0 \le s \le 2kc\, t_{\max }\),
Let \(k \ge 1\), and let \(A\) be a random \(k \times d\) real matrix on a probability space \((\Omega , \mu )\) whose \(kd\) entries \(A_{ij}\) are jointly independent, each Gaussian with law \(N(0, 1/k)\). Then for every fixed vector \(x \in \mathbb {R}^d\) and every \(\varepsilon \) with \(0 \lt \varepsilon \lt 1\),
where \(\| A(\omega )\, x\| \) denotes the Euclidean norm of the image \(A(\omega )\, x \in \mathbb {R}^k\).
Let \(0 \lt \varepsilon \lt 1\), let \(V\) be a finite set of points in \(\mathbb {R}^d\), and let \(f : \mathbb {R}^d \to \mathbb {R}^k\) be a linear map that is an \(\varepsilon \)-JL embedding of \(V\), so that for every pair \(u, v \in V\) one has \((1-\varepsilon )\, \| u-v\| ^2 \le \| f(u)-f(v)\| ^2 \le (1+\varepsilon )\, \| u-v\| ^2\). Then for all \(u, v \in V\),
Let \(\varepsilon \gt 0\) and let \(n \ge 2\) be an integer, and suppose the target dimension \(k\) satisfies \(k \ge 32 \log n / \varepsilon ^2\). Then for any finite set \(V\) of points in a Euclidean space with \(\left\lvert V\right\rvert \le n\), one has \(k \gt 0\) and
Let \((\Omega ,\mu )\) be a probability space, and let \(A\colon \Omega \to \mathbb {R}^{k\times d}\) be a measurable random \(k\times d\) matrix. Fix a finite set \(V\subseteq \mathbb {R}^d\) and a parameter \(\varepsilon \in \mathbb {R}\), and suppose that for every pair \(u,v\in V\) the single-vector bad event for the difference \(u-v\) has probability at most \(2\exp (-k\varepsilon ^2/8)\), that is,
Then the probability of the bad-pair event — that some ordered pair \((u,v)\in V\times V\) fails the two-sided distortion bound \((1-\varepsilon )\| u-v\| ^2 \le \| A(\omega )(u-v)\| ^2 \le (1+\varepsilon )\| u-v\| ^2\) — is at most
For \(\varepsilon \in \mathbb {R}\), vectors \(u, v \in \mathbb {R}^d\), and their images \(u', v' \in \mathbb {R}^k\), JLDistortion asserts the two-sided squared-distance inequality
Let \(\varepsilon \in \mathbb {R}\), let \(u, v \in \mathbb {R}^d\), and let \(A\) be a real \(k \times d\) matrix, acting as a linear map \(\mathbb {R}^d \to \mathbb {R}^k\). Suppose the difference \(u - v\) does not trigger the bad distortion event for \(A\); that is, the squared-norm distortion of \(u-v\) stays within the allowed band,
Then the images \(Au\) and \(Av\) satisfy the two-sided single-pair distortion bound
Let \(0 \lt \varepsilon \lt 1\) and let \(n \ge 2\) be an integer. Suppose the target dimension \(k\) satisfies \(k \ge 32\, \log n / \varepsilon ^2\). Then for every finite set \(V \subseteq \mathbb {R}^d\) of at most \(n\) points there exists a linear map \(f : \mathbb {R}^d \to \mathbb {R}^k\) that is an \(\varepsilon \)-JL embedding of \(V\); that is,
for all \(u, v \in V\).
Let \(\varepsilon \) be a real number with \(0 \lt \varepsilon \lt 1\), and let \(n \ge 2\) be an integer. Then there is a threshold \(k_0 \in \mathbb {N}\) such that for every target dimension \(k \ge k_0\), every ambient dimension \(d\), and every finite set \(V \subseteq \mathbb {R}^d\) of at most \(n\) points, there exists a linear map \(f : \mathbb {R}^d \to \mathbb {R}^k\) that is an \(\varepsilon \)-JL embedding of \(V\); that is, for all \(u, v \in V\),
Let \(\varepsilon \in (0,1)\) and let \(n \ge 2\) be an integer, and suppose the target dimension \(k\) satisfies \(k \ge 32\, \log n / \varepsilon ^2\). Then for every finite set \(V \subseteq \mathbb {R}^d\) of at most \(n\) points there exists a linear map \(f : \mathbb {R}^d \to \mathbb {R}^k\) such that, for all \(u, v \in V\),
where \(\lVert \cdot \rVert \) denotes the Euclidean norm.
Fix dimensions \(d\) and \(k\) with \(k \ge 1\), a distortion parameter \(\varepsilon \in (0,1)\), and a finite set of points \(V \subseteq \mathbb {R}^d\). Let \(A\) be a random \(k \times d\) matrix whose entries \(A_{ij}\) are independent Gaussian random variables, each with mean \(0\) and variance \(1/k\). If the union-bound quantity satisfies
then there exists a linear map \(f : \mathbb {R}^d \to \mathbb {R}^k\) that is an \(\varepsilon \)-JL embedding of \(V\): for every pair \(u, v \in V\),
Let \(k \ge 1\), and let \(A\) be a random \(k \times d\) real matrix defined on a probability space \((\Omega ,\mu )\), so that for each \(x \in \mathbb {R}^d\) the image \(Ax \in \mathbb {R}^k\) has coordinates \((Ax)_1,\dots ,(Ax)_k\). Suppose that for every \(x \in \mathbb {R}^d\) these \(k\) coordinates are independent, and that each coordinate \((Ax)_i\) is sub-Gaussian with variance proxy \(\| x\| ^2/k\), in the sense that \(t \mapsto e^{t(Ax)_i}\) is integrable and its moment generating function obeys \(\mathbb {E}\! \left[e^{t(Ax)_i}\right] \le \exp \! \bigl(\tfrac {\| x\| ^2}{k}\cdot \tfrac {t^2}{2}\bigr)\) for all \(t \in \mathbb {R}\), and moreover has second moment \(\mathbb {E}\! \left[(Ax)_i^2\right] = \| x\| ^2/k\). Fix \(0 \lt \varepsilon \lt 1\) and a finite set \(V \subseteq \mathbb {R}^d\) satisfying the union bound \(|V|^2 \cdot 2\exp (-k\varepsilon ^2/8) \lt 1\). Then there exists a linear map \(f : \mathbb {R}^d \to \mathbb {R}^k\) that is an \(\varepsilon \)-JL embedding of \(V\): for all \(u,v \in V\),
Let \(0 \lt \varepsilon \lt 1\), let \(n \ge 2\) be an integer, and let \(k\) be a target dimension satisfying \(k \ge 32\, \log n / \varepsilon ^2\). Then for every finite set \(V\) of at most \(n\) points in \(\mathbb {R}^d\) there exists a linear map \(f \colon \mathbb {R}^d \to \mathbb {R}^k\) such that every pair \(u, v \in V\) has its squared distance preserved up to a factor \(\varepsilon \):
Fix a distortion parameter \(\varepsilon \) with \(0 \lt \varepsilon \lt 1\) and an integer \(n \ge 2\). Then there is a threshold \(k_0 \in \mathbb {N}\) such that for every target dimension \(k \ge k_0\), every source dimension \(d\), and every finite set \(V \subseteq \mathbb {R}^d\) of at most \(n\) points, there exists a linear map \(f : \mathbb {R}^d \to \mathbb {R}^k\) that is an \(\varepsilon \)-JL embedding of \(V\): that is, \((1-\varepsilon )\, \| u - v\| ^2 \le \| f(u) - f(v)\| ^2 \le (1+\varepsilon )\, \| u - v\| ^2\) for all \(u, v \in V\).
Let \(\varepsilon \in (0,1)\) and let \(n \ge 2\) be an integer. Suppose \(k\) is a natural number with \(k \ge 32\, (\log n)/\varepsilon ^2\), and let \(V\) be a finite set of at most \(n\) points in \(\mathbb {R}^d\). Then there exists a linear map \(f : \mathbb {R}^d \to \mathbb {R}^k\) such that for all \(u, v \in V\),
A joint distribution \(\sigma \) on \(\mathrm{Fin}\, M \times \mathrm{Fin}\, N\) is a nonnegative function \(\sigma _{ij} \ge 0\) satisfying \(\sum _{i,j} \sigma _{ij} = 1\). It represents a correlated distribution over pure action profiles of the two players.
Let \(\sigma \) be a joint distribution on \(\mathrm{Fin}\, M \times \mathrm{Fin}\, N\), and let \(j\) be a column action. Then the column marginal \(\sigma ^{\mathrm{col}}_j = \sum _{i} \sigma _{ij}\) satisfies \(\sigma ^{\mathrm{col}}_j \ge 0\).
Let \(\sigma \) be a joint distribution on \(\mathrm{Fin}\, M \times \mathrm{Fin}\, N\), so \(\sigma _{ij} \ge 0\) for all \(i,j\) and \(\sum _{i,j} \sigma _{ij} = 1\). Then for every row action \(i\) the row marginal \(\sigma ^{\mathrm{row}}_i = \sum _{j} \sigma _{ij}\) satisfies \(\sigma ^{\mathrm{row}}_i \ge 0\).
Let \(\sigma \) be a joint distribution over \(\mathrm{Fin}\, M \times \mathrm{Fin}\, N\), and for each row action \(i\) let \(\sigma ^{\mathrm{row}}_i = \sum _{j} \sigma _{ij}\) denote its row marginal. Then the row marginals sum to one:
Let \(G\) be a loopless multigraph whose vertex set is finite, and suppose \(G\) has \(n \ge 4\) vertices. Then running Karger’s random-contraction algorithm \(\binom {n}{2} = n(n-1)/2\) times independently and returning the smallest cut found succeeds with probability
A run of the algorithm on \(G_0\) is indexed by a sequence of edge choices, one for each of the \(\left\lvert V(G_0)\right\rvert - 2\) contraction steps, and produces a multigraph. The current Lean definition is a placeholder that returns \(G_0\); the state-threading version is not yet formalized.
Given a multigraph \(G\) with at least one edge, a step of the algorithm is modelled as a function from an index \(i \in \mathrm{Fin}\, \left\lvert E(G)\right\rvert \) — thought of as a uniformly random edge choice — to the multigraph obtained by contracting the \(i\)-th edge of \(G\).
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\) of \(G\), and write \(n\) for the number of vertices of \(G\). Suppose that \(C\) is a minimum cut of \(G\), that \(C\) has positive size, and that \(n \ge 2\). Let \(p\) denote the probability that \(C\) survives \(n-2\) successive uniformly random edge-occurrence contractions. Then
and, moreover, every adaptive graph trace of \(n-2\) contraction steps whose initial graph is \(G\) ends in a multigraph with exactly two vertices.
Let \(n \ge 2\), and let \(T\) be an adaptive graph trace of \(n-2\) contraction steps whose initial multigraph \(T_0\) has exactly \(n\) vertices. Then the final multigraph \(T_{n-2}\), obtained after the \(n-2\) successive edge contractions prescribed by the trace, satisfies the termination condition: it has exactly two vertices.
Let \(S\) be a finite set, let \(p\) be a predicate on \(S\), and let \(f\colon S \to \mathbb {R}\). Suppose \(B \in \mathbb {R}\) is a lower bound for \(f\) on the elements satisfying \(p\), that is, \(B \le f(x)\) for every \(x \in S\) with \(p(x)\). Then
Let \(G\) be a multigraph on a vertex type \(\alpha \), and let \(e\) be an edge of \(G\), that is, a member of the edge multiset of \(G\). Then the occurrence \((e, 0)\) belongs to the graph-only occurrence space of \(G\); in other words, the pair consisting of \(e\) together with the index \(0\) is one of the enumerated edge occurrences of \(G\).
Let \(G\) be a multigraph and let \(C\) be a cut of \(G\), given by a vertex subset \(S\) that is nonempty and has nonempty complement. Then the size of \(C\) — the number of edges of \(G\) having exactly one endpoint in \(S\), counted with multiplicity — is at most the total number of edges of \(G\), again counted with multiplicity.
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(S\) of \(G\), and let \(\Omega \) be its sample space of edge occurrences — the finite set of pairs \((e, i)\) where \(e\) is an edge of \(G\) and \(i\) is an occurrence index distinguishing parallel copies of that edge. Call an occurrence *surviving* when its underlying edge \(e\) does not cross \(S\), that is, when it is not the case that one endpoint of \(e\) lies in \(S\) and the other lies outside \(S\). Then the number of surviving occurrences is the same whether they are counted directly among the elements of \(\Omega \) or among the elements of the equinumerous copy of \(\Omega \) obtained by regarding each member together with the fact that it belongs to \(\Omega \); in both counts the survival test is applied to the underlying occurrence, and the two cardinalities agree.
Let \(s\) be a state of the contraction process, with underlying multigraph \(G\), and let \((e, k)\) be any element of the sample space of edge occurrences of \(s\) — a pair consisting of an edge together with an index distinguishing parallel copies. Then the edge \(e\) is a member of the edge multiset of \(G\).
The finite set of pairs (edge, occurrence index) obtained from the edge multiset of \(s.G\); the second coordinate distinguishes parallel copies of the same edge, so that a uniform choice on this set is a uniform choice of an edge occurrence.
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a cut of \(G\), and let its sample space of edge occurrences be the finite set of pairs (edge, occurrence index) drawn from the edge multiset of \(G\). Then this sample space has cardinality equal to the edge count of \(G\), that is, the number of edges of \(G\) counted with multiplicity.
Let \(s\) be a state of the contraction process, consisting of a loopless multigraph \(G\) together with a distinguished cut \(S \subseteq V(G)\), and enumerate the edge multiset of \(G\) as a finite set of edge occurrences \((e, i)\), where the index \(i\) distinguishes parallel copies of the same edge. Say that an occurrence \((e, i)\) survives when its underlying edge \(e\) does not cross \(S\), that is, when it is not the case that one endpoint of \(e\) lies in \(S\) and the other lies outside \(S\). Then the number of surviving edge occurrences equals the number of edges of \(G\) that do not cross \(S\), counted with multiplicity.
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\) of \(G\). Form the survival ratio by enumerating the edge occurrences of \(G\) — the pairs of an edge with an index distinguishing its parallel copies — and dividing the number of occurrences whose underlying edge does not cross \(C\) by the total number of occurrences. Then this ratio equals the uniform survival probability of \(s\), namely the number of edges of \(G\) that do not cross \(C\), counted with multiplicity, divided by the total edge count of \(G\).
Let \(s\) be a state of the contraction process — a multigraph \(G\) together with a distinguished cut \(C\) of \(G\) — and write \(n\) for the number of vertices of \(G\). Suppose that \(C\) is a minimum cut of \(G\), that the size of \(C\) is strictly positive, and that \(n \ge 2\). Then the survival ratio of the edge-occurrence sample space of \(s\) is at least \(\tfrac {n-2}{n}\); that is,
where \(\rho (s)\) is the fraction of edge occurrences whose underlying edge does not cross \(C\).
Let \(s\) be a state consisting of a multigraph \(G\) on \(n = \left\lvert V(G)\right\rvert \) vertices together with a distinguished cut \(C\) of \(G\), and suppose that \(C\) is a minimum cut of \(G\), that its size is positive, and that \(n \ge 2\). Then the probability that \(C\) survives \(n-2\) successive uniformly random contractions of edge occurrences is at least
Let \(G\) be a multigraph on a vertex type \(\alpha \) with decidable equality, and let \(c\) be a sampled edge occurrence of \(G\), i.e. an element of its finite set of edge occurrences. Writing \(G_c\) for the multigraph obtained by contracting \(G\) along the edge underlying \(c\), the number of vertices of \(G_c\) equals the number of vertices of \(G\) minus one.
Let \(G\) be a multigraph, and let its graph-only occurrence space be the finite set of edge occurrences of \(G\), whose elements are pairs \((e, k)\) consisting of an edge \(e\) together with an index \(k\) distinguishing the copies of \(e\) among the parallel edges of \(G\). Then for every occurrence \((e, k)\) in this space, the edge \(e\) belongs to the edge multiset of \(G\).
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut of \(G\). Then the graph-only occurrence space built from \(G\) coincides, as a finite set of edge occurrences \((e, k) \in \mathrm{Sym}_2\, \alpha \times \mathbb {N}\), with the state’s sample space of edge occurrences; that is, both are the set of edge occurrences of the state’s underlying multigraph.
Let \(s\) be a state consisting of a multigraph \(G\) with \(n\) vertices together with a distinguished cut \(C\) of \(G\), and suppose that \(C\) is a minimum cut, that its size is strictly positive, and that \(n \ge 2\). Then two conclusions hold simultaneously. First, the recursive survival probability that \(C\) is preserved through the next \(n-2\) random edge-occurrence contractions of \(s\) is at least
Second, every adaptive graph trace of \(n-2\) contraction steps whose initial graph is \(G\) terminates in the output condition of the algorithm, that is, its graph after \(n-2\) steps has exactly two vertices.
Let \(G\) be a multigraph with at least two vertices, write \(n\) for its vertex count, and suppose \(H\) is obtained from \(G\) by a graph-only Karger run consisting of exactly \(n - 2\) successive edge-occurrence contractions. Then \(H\) satisfies the termination condition of the algorithm; that is, \(H\) has exactly two vertices.
Let \(G\) and \(H\) be multigraphs on a common vertex type, and suppose \(G\) has at least two vertices. If \(H\) is obtained from \(G\) by a graph-only Karger run consisting of exactly \(n - 2\) successive edge-occurrence contractions, where \(n\) is the number of vertices of \(G\), then \(H\) has exactly two vertices.
Let \(G\) and \(H\) be multigraphs on a vertex type \(\alpha \), and let \(\mathrm{steps} \in \mathbb {N}\). If \(H\) arises from \(G\) by a graph-only Karger run consisting of \(\mathrm{steps}\) successive edge-occurrence contractions, then the number of vertices of \(H\) equals the number of vertices of \(G\) minus \(\mathrm{steps}\), using truncated subtraction of natural numbers. Thus each contraction removes exactly one vertex, and after \(\mathrm{steps}\) of them the vertex count has dropped by \(\mathrm{steps}\).
An explicit indexed execution of the contraction loop: a family of graphs indexed by \(\mathbb {N}\), an edge occurrence of the graph at each index below \(\text{steps}\), and the requirement that the next graph is the contraction of the current one along that occurrence. The choice type depends on the current graph, matching the adaptive nature of the process.
Let \(n \ge 2\), and let \(T\) be an adaptive graph trace of \(n-2\) steps whose initial multigraph \(T_0\) has exactly \(n\) vertices. Then the multigraph \(T_{n-2}\) reached after all \(n-2\) contraction steps is a valid output of the algorithm; that is, it has exactly two vertices.
Let \(T\) be an adaptive graph trace of length \(\text{steps}\), so that \(T\) records a graph at each index, a sampled edge occurrence at each index below \(\text{steps}\), and produces each successive graph as the contraction of its predecessor along that occurrence. Then for every index \(i \le \text{steps}\), the graph-only Karger run relation holds for \(i\) steps between the initial graph \(T_0\) and the graph \(T_i\); that is, \(T_i\) is reachable from \(T_0\) by exactly \(i\) successive edge-occurrence contractions.
Let \((G_i)_{i \in \mathbb {N}}\) be an adaptive contraction trace over \(\text{steps}\) rounds on a vertex type with decidable equality: a family of multigraphs in which, for each index \(i \lt \text{steps}\), the graph \(G_{i+1}\) is obtained from \(G_i\) by contracting the edge underlying a sampled edge occurrence of \(G_i\). Then for every \(i \le \text{steps}\), the number of vertices of \(G_i\) equals the initial vertex count minus \(i\); that is, \(\lvert V(G_i)\rvert = \lvert V(G_0)\rvert - i\).
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\), whose two sides are given by a vertex subset \(S\) and its complement. Then the number of edges of \(G\) that do not cross \(C\) — those having no endpoint in \(S\) or no endpoint outside \(S\), counted with multiplicity — equals the total edge count of \(G\) minus the size of \(C\), the difference being taken as truncated subtraction in \(\mathbb {N}\):
Let \(s = (G, C)\) be a state of the contraction process, consisting of a multigraph \(G\) and a distinguished cut \(C\) with side \(S\), and let \((e, i)\) be an edge occurrence in the sample space of \(s\), so that \(e\) is an edge of \(G\) and \(i\) indexes its parallel copies. If this occurrence survives — that is, if \(e\) does not cross \(S\), meaning it is not the case that \(e\) has one endpoint in \(S\) and another endpoint outside \(S\) — then \(e\) does not belong to the multiset of crossing edges of \(C\).
Let \(s\) be a state of the contraction process, comprising a multigraph \(G\) together with a distinguished cut \(C\), and let \((e, i)\) be an edge occurrence of \(G\) drawn from its sample space of edge occurrences, whose underlying edge \(e\) does not cross \(C\). Then the multigraph carried by the successor state obtained by contracting \(s\) along this occurrence is equal to the multigraph obtained by contracting \(G\) itself along the same occurrence, namely the contraction \(G/e\).
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\) of \(G\), and let a sampled edge occurrence be given in the sample space of \(s\) — a pair (edge, occurrence index) drawn from the edge multiset of \(G\) — whose underlying edge does not cross \(C\). Then the pair formed by \(s\) and the successor state obtained by contracting \(s\) along that edge occurrence stands in the surviving-step relation; that is, the successor state arises from \(s\) by contracting some edge of \(G\) that does not cross \(C\).
The finite probability that the distinguished cut survives the next given number of random edge-occurrence choices, defined by recursion on the number of steps: zero steps have probability \(1\), and one further step averages over all edge occurrences, contributing \(0\) for occurrences that cross the cut and the recursive value at the contracted state otherwise.
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\) of \(G\), and write \(n\) for the number of vertices of \(G\). Suppose that \(C\) is a minimum cut of \(G\), that its size is positive, and that \(n \ge 2\). Then the probability that \(C\) survives \(n-2\) successive uniformly random edge-occurrence contractions is at least
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\), and let \(n\) be a natural number with \(2 \le n\). Suppose that \(C\) is a minimum cut of \(G\), that \(C\) has positive size, and that \(G\) has exactly \(n\) vertices. Then the recursive survival probability that \(C\) survives \(n-2\) successive uniformly random edge-occurrence contractions is at least
Let \(s\) be a state of the contraction process — a multigraph together with a distinguished cut — and assume that its sample space of edge occurrences is nonempty. Fix a natural number \(k\) and a real number \(B\), and suppose that for every edge occurrence that survives, meaning its underlying edge does not cross the distinguished cut, the recursive survival probability over \(k\) further choices, evaluated at the successor state obtained by contracting \(s\) along that occurrence, is at least \(B\). Then the recursive survival probability over \(k+1\) choices at \(s\) is at least the survival ratio of \(s\) multiplied by \(B\).
Let \(s\) be a state, consisting of a multigraph \(G\) together with a distinguished cut \(C\) of \(G\), and let \(\mathrm{steps}\) be a natural number. Write \(P_k(s)\) for the survival probability of the cut over \(k\) random edge-occurrence choices starting from \(s\). Then
where the sum ranges over the finite set of edge occurrences of \(G\) (each parallel copy of an edge counted separately), \(N\) is the number of such occurrences, and \(s_c\) denotes the state obtained by contracting \(s\) along the underlying edge of a surviving occurrence \(c\).
Let \(s\) and \(t\) be states of the contraction process, each consisting of a multigraph together with a distinguished cut, and suppose that \(t\) is obtained from \(s\) by a surviving run of length \(n\), that is, by \(n\) consecutive contractions of edges that do not cross the distinguished cut of the current state. Then the distinguished cut of \(t\) has the same size as the distinguished cut of \(s\), where the size of a cut is the number of its crossing edges counted with multiplicity.
Let \(s\) and \(t\) be states of the contraction process, each consisting of a multigraph together with a distinguished cut, and suppose that \(t\) is obtained from \(s\) by a surviving run of some length \(n\), that is, by \(n\) successive contractions of edges that do not cross the distinguished cut. If the distinguished cut \(C\) of \(s\) is a minimum cut of the multigraph \(s.G\), with size \(k\) (its number of crossing edges), then the number of vertices \(|V|\) and the number of edges \(|E|\) of the final multigraph \(t.G\) satisfy
Let \(s\) and \(t\) be states of the contraction process, each consisting of a multigraph together with a distinguished cut, and suppose \(t\) is reached from \(s\) by a surviving run of \(n\) steps. Write \(C\) for the distinguished cut of \(s\) and \(G\) for its multigraph, and let \(G'\) denote the multigraph of \(t\). If \(C\) is a minimum cut of \(G\), then
where \(\left\lvert C\right\rvert \) is the size of the cut (its number of crossing edges), \(\left\lvert V(G)\right\rvert \) is the number of vertices of \(G\), \(\left\lvert E(G')\right\rvert \) is the number of edges of \(G'\) counted with multiplicity, and the difference \(\left\lvert V(G)\right\rvert - n\) is truncated at \(0\).
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut of \(G\), and write \(n\) for the number of vertices of \(G\); suppose \(n \ge 2\). If a state \(t\) is reached from \(s\) by a surviving run of exactly \(n - 2\) consecutive surviving steps — each step contracting an edge of the current graph that does not cross the distinguished cut — then the multigraph of \(t\) satisfies the termination condition of the algorithm, that is, it has exactly two vertices.
Let \(s\) and \(t\) be states of the contraction process over a vertex type \(\alpha \), each consisting of a multigraph together with a distinguished cut of it. If \(t\) is reached from \(s\) by a surviving run of some number \(\mathrm{steps}\) of consecutive surviving steps — each contracting an edge that does not cross the current distinguished cut — then the underlying multigraph of \(t\) is reached from the underlying multigraph of \(s\) by a graph-only Karger run of the same length \(\mathrm{steps}\), that is, by \(\mathrm{steps}\) successive edge-occurrence contractions.
Let \(s\) and \(t\) be states of the contraction process, each consisting of a multigraph together with a distinguished cut of that multigraph, and suppose that \(t\) is obtained from \(s\) by a surviving run of some length \(n\), that is, by \(n\) consecutive contractions of edges that never cross the distinguished cut. If the distinguished cut of \(s\) is a minimum cut of its multigraph, then the distinguished cut of \(t\) is a minimum cut of its multigraph.
Let \(s\), \(t\), and \(u\) be states of the contraction process, each a multigraph paired with a distinguished cut. Suppose \(t\) is reached from \(s\) by a surviving run of \(i\) steps, that a further surviving step leads from \(t\) to \(u\), that the distinguished cut of \(s\) is a minimum cut of the multigraph of \(s\), and that the multigraph of \(t\) has at least two vertices. Writing \(n_t\) for the number of vertices of the multigraph of \(t\), \(c\) for the size of the distinguished cut of \(s\), and \(m\) for the number of edges of the multigraph of \(t\), then
Let \(s\) and \(t\) be states of the contraction process, each consisting of a multigraph together with a distinguished cut, and suppose that \(t\) is reached from \(s\) by a surviving run of length \(n\), that is, by \(n\) consecutive contractions of edges that do not cross the tracked cut. Writing \(G_s\) and \(G_t\) for the multigraphs of \(s\) and \(t\), and \(\lvert V(G)\rvert \) for the number of vertices of a multigraph \(G\), we have
where the subtraction is truncated subtraction on \(\mathbb {N}\) (equal to \(0\) once \(n\) exceeds \(\lvert V(G_s)\rvert \)).
Let \(s\) and \(t\) be states of the contraction process, each consisting of a multigraph together with a distinguished cut of that multigraph. If \(t\) arises from \(s\) by a surviving step — that is, \(t\) is obtained by contracting some edge \(e\) of the multigraph of \(s\) that does not cross the distinguished cut of \(s\) — then the two distinguished cuts have the same size; equivalently, the number of edges crossing the distinguished cut of \(t\), counted with multiplicity, equals the number of edges crossing the distinguished cut of \(s\).
Let \(s\) and \(t\) be states of the contraction process, each a multigraph equipped with a distinguished cut. If \(t\) arises from \(s\) by a surviving step — that is, by contracting some edge of the multigraph of \(s\) that does not cross the distinguished cut of \(s\) — then the multigraph of \(s\) has positive edge count.
Let \(s\) and \(t\) be states of the contraction process, each consisting of a multigraph together with a distinguished cut of it. If \(t\) arises from \(s\) by a surviving step—that is, by contracting some edge of \(s\)’s multigraph that does not cross its distinguished cut—then there is an edge occurrence in the graph-only occurrence space of \(s\)’s multigraph whose sampled contraction produces exactly the multigraph of \(t\).
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\) of \(G\), and let \(t\) be a state obtained from \(s\) by a surviving step: that is, \(t\) arises by contracting some edge \(e\) of \(G\) that does not cross \(C\), so that \(t\) has graph \(G/e\) and its distinguished cut is the cut induced by \(C\) on \(G/e\). If \(C\) is a minimum cut of \(G\), then the distinguished cut of \(t\) is a minimum cut of \(G/e\).
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\) of \(G\). Suppose that a surviving step is possible from \(s\), meaning that \(G\) has an edge that does not cross \(C\), whose contraction — merging the endpoints of that edge, deleting the resulting self-loops, and passing to the cut that \(C\) induces on the contracted graph — yields a further state of the process. Assume moreover that \(C\) is a minimum cut of \(G\) and that \(G\) has at least two vertices. Writing \(n\) for the number of vertices of \(G\), \(k\) for the size of \(C\) (its number of crossing edges, counted with multiplicity), and \(m\) for the number of edges of \(G\) (counted with multiplicity), one has
so that \(\frac{n-2}{n}\) is at most the one-step survival probability of \(C\).
Let \(s\) and \(t\) be states of the contraction process, each a multigraph together with a distinguished cut, and suppose that \(t\) arises from \(s\) by a surviving step, that is, \(t\) is obtained by contracting some edge \(e\) of the multigraph of \(s\) that does not cross the distinguished cut of \(s\). Then the number of vertices of the multigraph of \(t\) equals the number of vertices of the multigraph of \(s\) minus \(1\) (with the difference taken in \(\mathbb {N}\)).
Let \(T\) be an indexed surviving trace of \(\text{steps}\) steps, so that \(T\) assigns to each index \(i \in \mathbb {N}\) a state \(T_i\) and, for every \(i \lt \text{steps}\), the passage from \(T_i\) to \(T_{i+1}\) is a surviving step. Then
where \(r(T_i)\) denotes the occurrence-space survival ratio of \(T_i\) (the number of surviving edge occurrences divided by the total number of edge occurrences) and \(p(T_i)\) denotes its uniform one-step survival probability (the number of non-crossing edges divided by the total edge count).
Fix an integer \(n \ge 2\) and let \(T\) be an indexed surviving trace of length \(n-2\): a sequence of states \(T_0, T_1, \dots \) in which each \(T_{i+1}\) (for \(i \lt n-2\)) arises from \(T_i\) by contracting an edge of its multigraph that does not cross the distinguished cut. Suppose the distinguished cut of the initial state \(T_0\) is a minimum cut of the initial multigraph, that this initial multigraph has exactly \(n\) vertices, and that the multigraph of \(T_i\) has at least \(2\) vertices for every \(i \lt n-2\). Then the product over \(i \lt n-2\) of the one-step survival ratios \(\mathrm{surv}(T_i)\) — each the fraction of edge occurrences of \(T_i\) whose underlying edge does not cross the distinguished cut — satisfies
Let \(T\) be a surviving trace of length \(\text{steps}\), and for each index \(i\) write \(v_i\) for the number of vertices of the multigraph in its state at index \(i\). Suppose that \(\text{steps} = n - 2\) and that the initial graph has \(v_0 = n\) vertices. Then
where each factor is taken as a real number, the difference \(v_i - 2\) in the left-hand numerator being the truncated subtraction in \(\mathbb {N}\) and the right-hand differences being taken in \(\mathbb {R}\).
Let \(T\) be a surviving trace of \(\mathrm{steps}\) steps, with states \(s_0, s_1, \dots \); write \(G_i\) for the multigraph and \(C_i\) for the distinguished cut in state \(s_i\), and let \(n_i\) and \(m_i\) denote the vertex count and edge count of \(G_i\). Suppose the initial cut \(C_0\) is a minimum cut of \(G_0\), and that \(n_i \ge 2\) for every index \(i \lt \mathrm{steps}\). Then, writing \(c\) for the size of \(C_0\),
the right-hand factors being the one-step survival probabilities for cut size \(c\) and edge count \(m_i\).
Let \(T\) be a surviving trace of length \(\text{steps}\) over a vertex type \(\alpha \), that is, a family of states \((T_j)_{j \in \mathbb {N}}\) in which each consecutive pair \(T_j, T_{j+1}\) with \(j \lt \text{steps}\) is related by a surviving step. Then for every index \(i\) with \(i \le \text{steps}\), the states \(T_0, T_1, \dots , T_i\) constitute a surviving run of length \(i\) from the initial state \(T_0\) to the state \(T_i\).
Let \(T\) be an indexed surviving trace of \(n-2\) steps over a vertex type \(\alpha \), with states \(s_0, s_1, s_2, \ldots \), and suppose \(2 \le n\). Write \(c\) for the size of the distinguished cut of the initial state \(s_0\), and \(m_i\) for the edge count of the multigraph at state \(s_i\). Assume that the distinguished cut of \(s_0\) is a minimum cut of its multigraph, that the multigraph of \(s_0\) has exactly \(n\) vertices, and that for every \(i \lt n-2\) the multigraph of \(s_i\) has at least two vertices. Then
where each factor \(1 - c/m_i\) is the step-survival probability determined by the initial cut size \(c\) and the edge count \(m_i\).
Let \(T = (G_0, C_0), (G_1, C_1), \dots \) be a surviving trace: a family of states indexed by \(\mathbb {N}\) in which each state at an index below \(\text{steps}\) passes to its successor by contracting some edge of its current multigraph that does not cross its distinguished cut. Suppose the initial cut \(C_0\) is a minimum cut of \(G_0\), and fix an index \(i \lt \text{steps}\) at which the multigraph \(G_i\) has at least two vertices. Writing \(n_i\) for the number of vertices of \(G_i\), \(m_i\) for the number of edges of \(G_i\) (counted with multiplicity), and \(c\) for the size of the initial cut \(C_0\) (its number of crossing edges), one has
Let \(T\) be an indexed surviving trace of length \(\mathrm{steps}\), with states \(s_0, s_1, \dots \), and write \(c\) for the size of the distinguished cut of the initial state \(s_0\) and \(m_i\) for the edge count of the graph of state \(s_i\). Then the product of the uniform one-step survival probabilities of the states \(s_0, \dots , s_{\mathrm{steps}-1}\) equals
the product of the one-step survival probabilities computed with the initial cut size \(c\) at every index.
Let \(T\) be an indexed surviving trace of length \(\text{steps}\), recording at each index \(i\) a state consisting of a multigraph \(G_i\) together with a distinguished cut \(C_i\), with consecutive states below \(\text{steps}\) related by a surviving step. Then the product of the uniform one-step survival probabilities of the states of \(T\) equals the product, over all \(i \lt \text{steps}\), of the one-step survival probability \(1 - \dfrac {c_i}{m_i}\), where \(c_i\) is the size of the cut \(C_i\) (its number of crossing edges) and \(m_i\) is the total edge count of \(G_i\); that is,
Let \(n \ge 2\) and let \(T\) be a surviving trace of \(n-2\) steps, indexing states \(T_0, T_1, \ldots \) each consisting of a multigraph together with a distinguished cut. Suppose the distinguished cut of the initial state \(T_0\) is a minimum cut of its graph, that this initial graph has exactly \(n\) vertices, and that for every step index \(i \lt n-2\) the graph of \(T_i\) has at least two vertices. Then the product of the uniform one-step survival probabilities of the states \(T_0, \ldots , T_{n-3}\) satisfies
where \(p_i\) is the number of edges of \(T_i\) not crossing its distinguished cut divided by the total edge count of \(T_i\).
Let \(T\) be a surviving trace of length \(\mathrm{steps}\) over a vertex type \(\alpha \): a family of states \((G_j, C_j)\) indexed by \(j \in \mathbb {N}\), in which each consecutive pair below \(\mathrm{steps}\) is joined by a surviving step (the graph \(G_{j+1}\) arises from \(G_j\) by contracting an edge that does not cross the distinguished cut). Then for every index \(i \le \mathrm{steps}\), the number of vertices of the graph at index \(i\) equals the number of vertices of the initial graph at index \(0\) minus \(i\); that is, \(\lvert V(G_i)\rvert = \lvert V(G_0)\rvert - i\).
Let \(\alpha \) be a type with decidable equality, let \(m\) be a finite multiset over \(\alpha \), and let \(p\) be a decidable predicate on \(\alpha \). Form the enumeration set of \(m\), consisting of all pairs \((a, i)\) with \(a\) an element of \(m\) and \(0 \le i \lt \text{count}_m(a)\), so that each occurrence of each element of \(m\) is represented exactly once. Then the number of such pairs \((a,i)\) whose first coordinate \(a\) satisfies \(p\) equals the size (counted with multiplicity) of the sub-multiset of \(m\) consisting of the elements satisfying \(p\).
Let \(s\) be a state of the contraction process, consisting of a multigraph \(G\) together with a distinguished cut \(C\), and suppose \(G\) has at least one edge, so that its edge count \(m\) is positive. Then the uniform one-step survival probability of \(s\) — the number of non-crossing edge choices divided by \(m\) — equals
where \(c\) is the size of the cut \(C\), that is, its number of crossing edges (each counted with multiplicity).
Fix a prime \(p\) and let \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\). For a subset \(E \subseteq \{ 1,\dots ,n\} \), let \(r_E \colon V \to V_E\) be the restriction map that keeps the coordinates indexed by \(E\) and sets all others to zero. Then the kernel of \(r_E\) is the support submodule \(V_{E^c}\) consisting of those vectors whose support lies in the complement \(E^c = \{ 1,\dots ,n\} \setminus E\); that is, \(r_E\) annihilates exactly the vectors that vanish on every coordinate in \(E\).
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube with \(n\gt 0\) and \(\mathbb Ef=0\). Then \(I(f)\ge 1\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube. Then \(I(f)^2\le n\sum _iI_i(f)^2\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube and \(\varepsilon \gt 0\). There exist \(J\subseteq [n]\) and \(g:\{ 0,1\} ^n\to \mathbb R\) depending only on \(J\), with \(|J|\le 4nI(f)/\varepsilon \) and \(\mathbb E(f-g)^2\le \varepsilon \). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube with \(I(f)\gt 0\). Then \(\sum _i[I_i(f)/I(f)]\log (I(f)/I_i(f))\ge 0\), with zero influences contributing zero. For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube and \(\tau \gt 0\). Then \(|\{ i:I_i(f)\ge \tau \} |\le I(f)/\tau \). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube and \(\tau \in \mathbb R\). Define \(J_\tau (f)=\{ i\in [n]:I_i(f)\ge \tau \} \). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube with \(n\ge 2\) and \(\mathbb Ef=0\). Some coordinate \(i\) satisfies \(I_i(f)\ge \log n/(30n)\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube in dimension \(n\gt 0\). Then some \(i\in [n]\) satisfies \(I_i(f)\ge I(f)/n\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube, \(k\in \mathbb N\), and \(x\in \{ 0,1\} ^n\). Then \(f^{\le k}(x)+f^{\gt k}(x)=f(x)\), where the two parts retain Fourier coefficients with \(|S|\le k\) and \(|S|\gt k\), respectively.
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube and a positive integer \(k\). Then \(\mathbb E(f-f^{\le k})^2\le I(f)/k\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube, \(k\in \mathbb N\), and \(\tau \gt 0\). There exists \(g:\{ 0,1\} ^n\to \mathbb R\) depending only on \(J_\tau (f)=\{ i:I_i(f)\ge \tau \} \) such that \(\mathbb E(f^{\le k}-g)^2\le n\tau \). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube with \(I(f)\gt 0\). There exists \(i\in [n]\) with \(I_i(f)\ge \sum _jI_j(f)^2/I(f)\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube, \(i\in [n]\), and \(0\le \rho \le 1\). Then \(I_i^{(\rho )}(f)\le I_i(f)\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube and \(i\in [n]\). Then \(I_i^{(1)}(f)=I_i(f)\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \{ -1,1\} \) be a function on the uniform Boolean cube, \(i\in [n]\), and \(0\lt \rho \le 1\). Then \(I_i^{(\rho )}(f)\le I_i(f)\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation.
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube and a positive integer \(k\). Then \(\sum _{|S|\gt k}\widehat f(S)^2\le I(f)/k\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
Let \(f:\{ 0,1\} ^n\to \mathbb R\) be a function on the uniform Boolean cube. Then \(I(f)=\sum _{i=0}^{n-1}I_i(f)\). For a real cube function \(h\), define \(I_i(h)=\tfrac 14\mathbb E_x(h(x)-h(x^{(i)}))^2\), where \(x^{(i)}\) flips coordinate \(i\) and \(\mathbb E_x\) is uniform expectation, and \(I(h)=\sum _{i=0}^{n-1}I_i(h)\).
A subspace \(C \le \mathcal{H}_n\) satisfies the Knill–Laflamme condition for \(t\)-error correction if for all Pauli strings \(E,F\) with \(\mathrm{wt}(E),\mathrm{wt}(F) \le t\) there exists \(\lambda _{EF}\in \mathbb {C}\) such that \(P_C\, E^\dagger F\, P_C = \lambda _{EF}\, P_C\), where \(P_C\) is the orthogonal projection onto \(C\).
For naturals \(n\), \(j\), \(x\),
where the subtraction \(n-x\) is natural-number subtraction; the definition gives the intended values for \(x \le n\).
Let \(n\) and \(x\) be natural numbers with \(x \le n\), and let \(K_j^{(n)}(x)\) denote the binary Krawtchouk polynomial. Then for every real \(z\),
Fix a natural number \(n\), and let \(r, s\) be natural numbers with \(r \le n\) and \(s \le n\). Then the binary Krawtchouk polynomials satisfy
that is, they are orthogonal with respect to the weights \(\binom {n}{x}\) on \(\{ 0, 1, \dots , n\} \).
Let \(n\) and \(t\) be natural numbers with \(t + 1 \le n\), and let \(K_j^{(n)}\) denote the degree-\(j\) Krawtchouk polynomial associated with \(n\). Suppose \(a\) is a real number lying strictly below every real zero of \(K_t^{(n)}\): that is, \(a \lt \xi \) whenever \(K_t^{(n)}(\xi ) = 0\). Then \(K_j^{(n)}(a) \gt 0\) for every \(j\) with \(j \le t\).
Let \(n\), \(j\), and \(x\) be natural numbers with \(1 \le j\), \(j + 1 \le n\), and \(x \le n\). Then the binary Krawtchouk polynomials \(K_j^{(n)}\) satisfy the three-term recurrence
Fix a natural number \(n\), and let \(j, x\) be natural numbers with \(j \le n\) and \(x \le n\). Then the Krawtchouk polynomial \(K_j^{(n)}\), evaluated at the real number \(x\), equals the alternating sum
Let \(l\) be a finite list of weighted edges over \(n\) vertices, and let \(e\) be a weighted edge that occurs in \(l\). Then \(l\) is a permutation of the list obtained by deleting the first occurrence of \(e\) from \(l\) and appending \(e\) at the end.
Let \(\mathit{base}\) and \(\mathit{rest}\) be lists of weighted edges on \(n\) vertices, and let \(g\) and \(e\) be two further weighted edges, with respective endpoints \(u_g, v_g\) and \(u_e, v_e\). Suppose that, using only the edges in the concatenation \(\mathit{base} \mathbin {+\! \! +} \mathit{rest}\), the vertex \(u_e\) is connected to \(u_g\) and the vertex \(v_g\) is connected to \(v_e\). Then for every pair of vertices \(a, b\) that are connected using the edges of \(\mathit{base} \mathbin {+\! \! +} (g :: \mathit{rest})\), the vertices \(a\) and \(b\) are also connected using the edges of \(\mathit{base} \mathbin {+\! \! +} \mathit{rest} \mathbin {+\! \! +} [e]\).
Fix \(n\) and consider weighted edges on a vertex set of size \(n\), each edge carrying a first and a second endpoint. For a list of edges \(L\), write \(a \leadsto _L b\) to mean that \(a\) reaches \(b\) through \(L\), i.e. \(a\) and \(b\) are joined by a path whose edges lie in \(L\) (the reflexive–transitive closure of the symmetric adjacency relation induced by \(L\)). Let \(\mathit{base}\) and \(\mathit{rest}\) be lists of weighted edges, and let \(g\) and \(e\) be weighted edges, with first and second endpoints \(g_1, g_2\) and \(e_1, e_2\) respectively. Suppose that, using the edges in the concatenation \(\mathit{base} \mathbin {+\! \! +} \mathit{rest}\), the endpoint \(e_1\) reaches \(g_2\) and the endpoint \(g_1\) reaches \(e_2\). Then for every pair of vertices \(a, b\): if \(a \leadsto b\) using the list obtained by placing \(g\) at the front of \(\mathit{rest}\) and appending that to \(\mathit{base}\), then \(a \leadsto b\) using \(\mathit{base} \mathbin {+\! \! +} \mathit{rest} \mathbin {+\! \! +} [e]\), the concatenation of \(\mathit{base}\), \(\mathit{rest}\), and the single edge \(e\).
Let \(\mathit{base}\) and \(S\) be lists of weighted edges on \(n\) vertices, and let \(e\) be a weighted edge whose two endpoints are distinct. Suppose the endpoints of \(e\) are connected by a path using the edges of \(\mathit{base}\) together with those of \(S\), but are not connected using the edges of \(\mathit{base}\) alone. Then there is an edge \(f\) occurring in \(S\) such that, for every pair of vertices \(a\) and \(b\) connected by a path through the edges of \(\mathit{base}\) together with \(S\), the vertices \(a\) and \(b\) are still connected by a path through the edges of \(\mathit{base}\) together with the edges of \(S\) after removing one occurrence of \(f\) and adjoining \(e\).
Let \(E\) be a list of weighted edges on \(n\) vertices, and let \(S\) be another such list in which every edge is also an edge of \(E\). Suppose that \(S\) spans like \(E\): whenever two vertices \(u\) and \(v\) are joined by a path through the edges of \(E\), they are also joined by a path through the edges of \(S\). Let \(K\) denote the edge list produced by Kruskal’s algorithm on \(E\), and write \(w(e)\) for the weight of an edge \(e\). Then the total weight of \(K\) is no greater than that of \(S\):
Fix \(n \in \mathbb {N}\), and let \(E\) be a list of weighted edges on the vertex set \(\mathrm{Fin}\, n\). Then the edge list \(\mathrm{kruskal}\, n\, E\) produced by Kruskal’s algorithm on \(E\) spans \(E\): for every pair of vertices \(u, v \in \mathrm{Fin}\, n\), if \(u\) and \(v\) are connected by a path through the edges of \(E\), then they are also connected by a path through the selected edges \(\mathrm{kruskal}\, n\, E\).
Fix \(n \in \mathbb {N}\) and a list of weighted edges over \(n\) vertices. Every edge returned by running Kruskal’s algorithm on this list occurs in the original list: if \(e\) belongs to the output of \(\mathtt{kruskal}\, n\, \mathit{edges}\), then \(e\) belongs to \(\mathit{edges}\).
Let \(\mathit{uf}\colon \mathrm{Fin}\, n \to \mathbb {N}\) be a union-find state on \(n\) nodes, and let \(i,j\) be two nodes whose representatives agree, \(\mathit{uf}(i) = \mathit{uf}(j)\). Then merging the components of \(i\) and \(j\) has no effect: the merged state equals \(\mathit{uf}\).
Let \(n\) be a natural number, let \(\mathit{uf}\colon \mathrm{Fin}\, n \to \mathbb {N}\) be a union-find state, and let \(a\) and \(b\) be weighted edges over \(n\) vertices, with endpoints \(a_u, a_v\) and \(b_u, b_v\) respectively. Form one state by merging the components of \(a_u\) and \(a_v\) and then merging the components of \(b_u\) and \(b_v\), and a second state by performing these two merges in the opposite order. Then the two resulting states induce the same partition of the nodes: for all \(i, j : \mathrm{Fin}\, n\), the first state assigns \(i\) and \(j\) the same representative if and only if the second state does.
Fix \(n \in \mathbb {N}\) and a union-find state \(\mathit{uf} \colon \mathrm{Fin}\, n \to \mathbb {N}\) on \(n\) nodes. For any list \(l\) of weighted edges over \(n\) vertices and any single weighted edge \(e\) over \(n\) vertices with endpoints \(u\) and \(v\), merging all edges of the list \(l \mathbin {+\! \! +} [e]\) into \(\mathit{uf}\) in order produces the same union-find state as first merging all edges of \(l\) into \(\mathit{uf}\) in order and then merging the components of \(u\) and \(v\) in the state so obtained.
Let \(\mathit{uf}_1, \mathit{uf}_2 : \mathrm{Fin}\, n \to \mathbb {N}\) be two union-find states over \(n\) nodes that induce the same partition, meaning that for all nodes \(i, j\) one has \(\mathit{uf}_1(i) = \mathit{uf}_1(j)\) if and only if \(\mathit{uf}_2(i) = \mathit{uf}_2(j)\). Then for any list of weighted edges, folding that list into each state by successively merging the endpoints of every edge yields two states that again induce the same partition.
Let \(\mathit{uf} : \mathrm{Fin}\, n \to \mathbb {N}\) be a union-find state on \(n\) nodes, let \(e\) be a weighted edge with endpoints \(u\) and \(v\) over these nodes, and let \(S\) be a list of weighted edges over the same nodes. Folding the list \(e :: S\) into \(\mathit{uf}\) and folding the list \(S\) into the union-find state that merges the components of \(u\) and \(v\) in \(\mathit{uf}\) induce the same partition of the nodes: for all nodes \(i\) and \(j\), folding \(e :: S\) into \(\mathit{uf}\) assigns \(i\) and \(j\) a common representative if and only if folding \(S\) into the merge of \(u\) and \(v\) assigns \(i\) and \(j\) a common representative.
Fix \(n\) and let \(uf : \mathrm{Fin}\, n \to \mathbb {N}\) be a union-find state assigning each of the \(n\) nodes a representative. Let \(\mathit{base}\) and \(\mathit{edges}\) be lists of weighted edges on these \(n\) vertices, and suppose \(uf\) already encodes reachability along \(\mathit{base}\), in the sense that for all nodes \(a, b\) one has \(uf(a) = uf(b)\) if and only if \(a\) and \(b\) are connected by a path through the edges of \(\mathit{base}\). Then, writing \(uf'\) for the state obtained by merging the endpoints of every edge of \(\mathit{edges}\) into \(uf\) in order, for all nodes \(a, b\) one has \(uf'(a) = uf'(b)\) if and only if \(a\) and \(b\) are connected by a path through the edges of the concatenated list \(\mathit{base} \mathbin {+\! \! +} \mathit{edges}\).
Fix a natural number \(n\), and let \(\mathit{edges}\) be a list of weighted edges on the vertex set \(\mathrm{Fin}\, n\). Consider the union-find state obtained by starting from the initial state on \(n\) nodes, in which every node is its own representative, and processing \(\mathit{edges}\) in order, so that at each weighted edge the components of the two endpoints of that edge are merged. Then, for any two vertices \(a, b \in \mathrm{Fin}\, n\), the vertices \(a\) and \(b\) carry the same representative in the resulting state if and only if \(a\) and \(b\) are connected by a path through the edges of \(\mathit{edges}\) — equivalently, if and only if the reflexive–transitive closure of symmetric adjacency along \(\mathit{edges}\) relates \(a\) to \(b\).
Fix \(n\) and an initial union-find state \(\mathit{uf} : \mathrm{Fin}\, n \to \mathbb {N}\). If two lists of weighted edges \(l_1\) and \(l_2\) over \(n\) vertices are permutations of one another, then folding each list into \(\mathit{uf}\) by merging the endpoints of every edge in turn yields two union-find states that induce the same partition of the nodes; that is, for all \(i, j : \mathrm{Fin}\, n\), the two states assign \(i\) and \(j\) equal representatives under one exactly when they do under the other.
Let \(\mathit{uf} \colon \mathrm{Fin}\, n \to \mathbb {N}\) be a union-find state assigning each of \(n\) nodes a representative, and let \(\mathit{es}\) be a list of weighted edges over these \(n\) nodes. Suppose two nodes \(a, b \in \mathrm{Fin}\, n\) satisfy \(\mathit{uf}(a) = \mathit{uf}(b)\). Folding every edge of \(\mathit{es}\) into \(\mathit{uf}\) in sequence yields a union-find state \(\mathit{uf}'\) for which \(\mathit{uf}'(a) = \mathit{uf}'(b)\); that is, sequential merging never separates two nodes that already carry a common representative.
\(\mathtt{processEdges}\, \mathit{es}\, \mathit{uf}\, \mathit{acc}\) scans the edge list \(\mathit{es}\) left-to-right: an edge \(e\) is added to the accumulator and its endpoints are merged whenever \(\mathit{uf}.\mathtt{find}\, e.u \ne \mathit{uf}.\mathtt{find}\, e.v\); otherwise \(e\) is skipped. The reversed accumulator is returned when the list is exhausted.
Fix \(n \in \mathbb {N}\), and let \(\mathit{es}\) be a list of weighted edges over \(n\) vertices, \(\mathit{uf}\) a union-find state on \(n\) nodes, and \(\mathit{acc}\) a list of weighted edges over \(n\) vertices. Then running the edge-selection sweep on \(\mathit{es}\) from state \(\mathit{uf}\) with initial accumulator \(\mathit{acc}\) yields the reversal of \(\mathit{acc}\) concatenated (on the left) with the result of running the same sweep on \(\mathit{es}\) from the same state \(\mathit{uf}\) but with the empty accumulator.
Fix \(n\), and let \(\mathit{es}\) be a list of weighted edges over \(n\) vertices, let \(\mathit{uf}\) be a union-find state on \(n\) nodes, and let \(\mathit{acc}\) be an accumulator list of weighted edges. If a weighted edge \(e\) occurs in the output produced by running the edge-selection sweep on the edge list \(\mathit{es}\), the union-find state \(\mathit{uf}\), and the accumulator \(\mathit{acc}\), then \(e\) already occurs in the accumulator \(\mathit{acc}\) or \(e\) occurs in the edge list \(\mathit{es}\).
Let \(\mathit{edges}\) be a list of weighted edges on \(n\) vertices, sorted in nondecreasing order of weight, and let \(\mathit{uf}\) be a union-find state whose induced partition is exactly the reachability relation of some base edge list \(\mathit{base}\): for all vertices \(a\) and \(b\), the representatives \(\mathit{uf}(a)\) and \(\mathit{uf}(b)\) are equal if and only if \(a\) and \(b\) are joined by a path through the edges of \(\mathit{base}\). Let \(S\) be any list of edges, each of which occurs in \(\mathit{edges}\), such that merging all of \(S\) into \(\mathit{uf}\) and merging all of \(\mathit{edges}\) into \(\mathit{uf}\) yield the same partition of the vertices. Then the greedy sweep — which starts from \(\mathit{uf}\), scans \(\mathit{edges}\) in order, and keeps each edge exactly when its two endpoints lie in currently distinct components — produces a set of edges whose total weight is at most the total weight of \(S\).
Let \(n \in \mathbb {N}\), let \(\mathit{uf} : \mathrm{Fin}\, n \to \mathbb {N}\) be a union-find state, and let \(\mathit{es}\) be a list of weighted edges over \(n\) vertices. Merging into \(\mathit{uf}\) only the edges retained by the edge-selection sweep (running the sweep from \(\mathit{uf}\) with an empty accumulator) yields a union-find state inducing the same equivalence relation on nodes as merging all of \(\mathit{es}\) into \(\mathit{uf}\): for all \(i, j : \mathrm{Fin}\, n\), the two nodes share a representative after the first process exactly when they share a representative after the second.
Fix a number of nodes \(n\), a union-find state \(\mathit{uf} : \mathrm{Fin}\, n \to \mathbb {N}\), a weighted edge \(e\) over \(n\) vertices with endpoints \(u\) and \(v\), and a list \(\mathit{rest}\) of further weighted edges over \(n\) vertices. Suppose the two endpoints of \(e\) already have the same representative in \(\mathit{uf}\), that is, \(\mathit{uf}(u) = \mathit{uf}(v)\). Then the edge-selection sweep applied to the list \(e :: \mathit{rest}\) starting from state \(\mathit{uf}\) with empty accumulator returns exactly the same edge list as the sweep applied to \(\mathit{rest}\) starting from the same state \(\mathit{uf}\) with empty accumulator.
Fix a natural number \(n\) and a union-find state \(\mathit{uf}\) over \(n\) nodes, and let \(e\) be a weighted edge over \(n\) vertices — with endpoints \(u_e, v_e\) and weight \(w_e\) — followed by a list \(\mathit{rest}\) of further weighted edges. Suppose the endpoints of \(e\) lie in different components, that is, \(\mathit{uf}(u_e) \ne \mathit{uf}(v_e)\). Then the total weight of the edge-selection sweep applied to the list \(e :: \mathit{rest}\) starting from \(\mathit{uf}\) and an empty accumulator equals \(w_e\) plus the total weight of the sweep applied to \(\mathit{rest}\) starting from an empty accumulator and the state obtained by merging the components of \(u_e\) and \(v_e\) in \(\mathit{uf}\).
Let \(g\) be a weighted edge on \(n\) vertices with endpoints \(p\) and \(q\), let \(\mathit{rest}\) be a list of weighted edges on the same \(n\) vertices, and let \(a\) and \(b\) be vertices; say two vertices are \(\mathit{rest}\)-connected when they are joined by a path using only the edges in \(\mathit{rest}\). Then \(a\) and \(b\) are joined by a path through the edges of the list obtained by prepending \(g\) to \(\mathit{rest}\) if and only if at least one of the following holds: \(a\) and \(b\) are \(\mathit{rest}\)-connected; or \(a\) is \(\mathit{rest}\)-connected to \(p\) while \(q\) is \(\mathit{rest}\)-connected to \(b\); or \(a\) is \(\mathit{rest}\)-connected to \(q\) while \(p\) is \(\mathit{rest}\)-connected to \(b\).
Let \(\mathit{edges}\) and \(\mathit{edges}'\) be two lists of weighted edges on \(n\) vertices, and suppose that whenever two vertices \(a\) and \(b\) are symmetrically adjacent in \(\mathit{edges}\) — that is, some edge of \(\mathit{edges}\) has endpoints \(\{ a,b\} \) — the vertices \(a\) and \(b\) are connected by a path through \(\mathit{edges}'\). Then for any vertices \(u\) and \(v\), if \(u\) and \(v\) are connected by a path through \(\mathit{edges}\), they are also connected by a path through \(\mathit{edges}'\).
Fix \(n \in \mathbb {N}\), and let \(l_1\) and \(l_2\) be two lists of weighted edges on the vertex set \(\mathrm{Fin}\, n\) that have the same members, i.e. \(x \in l_1 \Leftrightarrow x \in l_2\) for every weighted edge \(x\). Then for any vertices \(a, b \in \mathrm{Fin}\, n\), the vertex \(b\) is reachable from \(a\) through the edges of \(l_1\) if and only if it is reachable from \(a\) through the edges of \(l_2\).
Fix \(n\), let \(e_1\) and \(e_2\) be lists of weighted edges on the vertex set \(\mathrm{Fin}\, n\), and let \(u, v : \mathrm{Fin}\, n\). Suppose \(u\) and \(v\) are connected by a path through the edges of \(e_1\), and that every edge of \(e_1\) also occurs in \(e_2\). Then \(u\) and \(v\) are connected by a path through the edges of \(e_2\).
Let \(\mathit{edges}\) be a list of weighted edges on \(n\) vertices, and let \(e\) be a weighted edge belonging to \(\mathit{edges}\). Then the endpoints \(u\) and \(v\) of \(e\) satisfy \(\mathtt{Reach}\, \mathit{edges}\, u\, v\); that is, they are connected by a path through \(\mathit{edges}\).
Let \(\mathit{edges}\) be a list of weighted edges on \(n\) vertices, and call two vertices reachable when they are joined by a path along these edges (in the sense of the reflexive–transitive closure of symmetric adjacency). Then reachability is symmetric: for all vertices \(u, v\), if \(u\) is reachable from \(v\), then \(v\) is reachable from \(u\).
Let \(\mathit{uf}\) be a union-find state on \(n\) nodes, let \(e\) be a weighted edge with endpoints \(u\) and \(v\), and let \(\mathit{rest}\) and \(S\) be lists of weighted edges. Suppose every edge of \(S\) occurs in the list \(e :: \mathit{rest}\) formed by prepending \(e\) to \(\mathit{rest}\); that merging all edges of \(S\) into \(\mathit{uf}\) induces the same partition of the nodes as merging all edges of \(\mathit{rest}\) into \(\mathit{uf}\); and that \(e\) is already redundant for \(\mathit{uf}\), meaning \(\mathit{uf}(u) = \mathit{uf}(v)\). Then there exists a list \(S'\) of weighted edges, every one of which occurs in \(\mathit{rest}\), such that merging all edges of \(S'\) into \(\mathit{uf}\) induces the same partition as merging all edges of \(\mathit{rest}\), and the total weight of \(S'\) is at most the total weight of \(S\).
Given a list of weighted edges \(\mathit{edges}\) on \(n\) vertices, \(\mathtt{SymAdj}\, \mathit{edges}\, u\, v\) holds when there exists an edge \(e \in \mathit{edges}\) whose endpoints are \(\{ u, v\} \) (in either order), i.e. \((e.u = u \wedge e.v = v) \vee (e.u = v \wedge e.v = u)\).
Fix \(n \in \mathbb {N}\) and two vertices \(u, v \in \mathrm{Fin}\, n\), and let \(e_1\) and \(e_2\) be lists of weighted edges on \(n\) vertices such that every edge of \(e_1\) also occurs in \(e_2\). If \(u\) and \(v\) are symmetrically adjacent in \(e_1\) — that is, some edge of \(e_1\) has endpoints \(\{ u,v\} \) in either order — then \(u\) and \(v\) are symmetrically adjacent in \(e_2\).
Fix a list of weighted edges on \(n\) vertices, and let \(u\) and \(v\) be two vertices. Recall that \(u\) and \(v\) are symmetrically adjacent when some edge in the list has \(\{ u, v\} \) as its pair of endpoints, in either order. If \(u\) and \(v\) are symmetrically adjacent, then so are \(v\) and \(u\).
Let \(\mathit{edges}\) be a list of weighted edges over \(n\) vertices, and let \(e\) be a weighted edge occurring in \(\mathit{edges}\). Then the total weight of \(\mathit{edges}\) equals the weight of \(e\) plus the total weight of the list obtained by deleting the first occurrence of \(e\) from \(\mathit{edges}\).
Let \(\mathit{uf} \colon \mathrm{Fin}\, n \to \mathbb {N}\) be a union-find state on \(n\) nodes, and let \(i\) be one of its nodes. Then the representative of \(i\) returned by \(\mathrm{find}\) coincides with the direct evaluation of the state at \(i\); that is, \(\mathrm{find}(\mathit{uf}, i) = \mathit{uf}(i)\).
Consider the identity union-find on \(n\) nodes, in which every node is its own representative. For every node \(i \in \mathrm{Fin}\, n\), applying the find operation to this initial state returns the underlying numeric value of \(i\); that is, the representative of \(i\) is \(i\) itself.
\(\mathtt{SamePartition}\, \mathit{uf}_1\, \mathit{uf}_2\) holds when two union-find states induce the same equivalence relation on nodes: for all \(i, j : \mathrm{Fin}\, n\), \(\mathit{uf}_1(i) = \mathit{uf}_1(j)\) if and only if \(\mathit{uf}_2(i) = \mathit{uf}_2(j)\).
Fix \(n \in \mathbb {N}\), and recall that a union-find state on \(n\) nodes is a function \(\mathrm{Fin}\, n \to \mathbb {N}\) assigning to each node a representative label. Every such state \(\mathit{uf}\) induces the same partition as itself: for all nodes \(i, j : \mathrm{Fin}\, n\), one has \(\mathit{uf}(i) = \mathit{uf}(j)\) if and only if \(\mathit{uf}(i) = \mathit{uf}(j)\).
Fix \(n \in \mathbb {N}\), and regard a union-find state on \(n\) nodes as a function \(\mathrm{Fin}\, n \to \mathbb {N}\) assigning each node a representative. If two such states \(\mathit{uf}_1\) and \(\mathit{uf}_2\) induce the same partition — that is, for all nodes \(i, j\) one has \(\mathit{uf}_1(i) = \mathit{uf}_1(j)\) if and only if \(\mathit{uf}_2(i) = \mathit{uf}_2(j)\) — then \(\mathit{uf}_2\) and \(\mathit{uf}_1\) likewise induce the same partition.
Let \(n\) be a natural number, and let \(\mathit{uf}_1, \mathit{uf}_2, \mathit{uf}_3 : \mathrm{Fin}\, n \to \mathbb {N}\) be three union-find states on \(n\) nodes. Suppose that \(\mathit{uf}_1\) and \(\mathit{uf}_2\) induce the same partition of the nodes, and that \(\mathit{uf}_2\) and \(\mathit{uf}_3\) induce the same partition. Then \(\mathit{uf}_1\) and \(\mathit{uf}_3\) induce the same partition: for all nodes \(i, j\), one has \(\mathit{uf}_1(i) = \mathit{uf}_1(j)\) if and only if \(\mathit{uf}_3(i) = \mathit{uf}_3(j)\).
Let \(\mathit{uf}\colon \mathrm{Fin}\, n \to \mathbb {N}\) be a union-find state whose classes coincide with the reachability classes of an edge list \(\mathit{base}\): for all nodes \(a,b\) one has \(\mathit{uf}(a) = \mathit{uf}(b)\) if and only if \(a\) and \(b\) are joined by a path through the edges of \(\mathit{base}\). Let \(S\) be a further list of weighted edges, and let \(e\) be a weighted edge with endpoints \(u\) and \(v\) such that merging every edge of \(S\) into \(\mathit{uf}\) places \(u\) and \(v\) in a common component, while \(\mathit{uf}\) itself keeps them in distinct components. Then there is an edge \(f \in S\) for which the state obtained by first merging \(u\) with \(v\) and then merging all edges of \(S\) other than \(f\) induces the same partition of the nodes as the state obtained by merging all edges of \(S\) into \(\mathit{uf}\).
The union–find structure obtained by running the greedy loop over a list of weighted edges: starting from a state \(uf\), each edge \(e\) is merged into the state when its endpoints \(e.u\) and \(e.v\) currently lie in different components, and is ignored otherwise. Unlike \(\texttt{Kruskal.processEdges}\) it returns only the final state, not the list of accepted edges.
Fix \(n \in \mathbb {N}\), let \(\mathit{es}\) be a list of weighted edges over \(n\) vertices, and let \(\mathit{uf} : \mathrm{Fin}\, n \to \mathbb {N}\) be an initial union–find state. The greedy loop, which scans \(\mathit{es}\) in order and merges the endpoints of each edge only when they currently carry different representatives, yields a final state that induces the same partition of the vertices as the unconditional fold, which merges the endpoints of every edge of \(\mathit{es}\) in order without regard to their current components. That is, for all \(i, j : \mathrm{Fin}\, n\), nodes \(i\) and \(j\) share a representative in the greedy-loop state if and only if they share a representative in the unconditional-fold state.
A \(\texttt{LearnerPolicy}\ N\) is a deterministic online strategy for a learner interacting with \(N\) experts: it provides a weight function \(\mathrm{weights} : (t : \mathbb {N}) \to \texttt{LossHistory}\ N\ t \to \mathrm{Fin}\, N \to \mathbb {R}\) together with proofs that the weights are nonnegative and sum to one, so that the weight vector at each round forms a valid mixed strategy.
Let \(\alpha \) be a linearly ordered type, let \(x \in \alpha \), and let \(xs\) be a finite list of elements of \(\alpha \). The return value of the timed ordered insertion of \(x\) into \(xs\) is a list of length \(\left\lvert xs\right\rvert + 1\).
Let \(l\) be a literal over \(n\) Boolean variables, given by a variable index and a polarity that marks it as positive or negated, and let \(x \in \{ \text{true},\text{false}\} ^n\) be an assignment. Writing \(\bar{l}\) for the literal obtained from \(l\) by flipping its polarity while keeping the same variable, the value of \(\bar{l}\) under \(x\) is the negation of the value of \(l\) under \(x\); that is, \(\bar{l}(x) = \lnot \, l(x)\).
Fix \(n \in \mathbb {N}\), and recall that a literal is a pair consisting of a variable in \(\{ 0, 1, \dots , n-1\} \) together with a polarity bit, and that flipping a literal’s polarity toggles this bit while leaving the variable unchanged. Then the polarity-flipping operation is injective: if two literals have the same polarity flip, they are equal.
Let \(n, l \in \mathbb {N}\) with \(l \ge 1\), and let \(\mathcal{D}\) be a layer-2 datum on \(n\) variables whose gates \(g_0, \dots , g_{m-1}\) are DNF formulas. Let \(C\) be a Boolean circuit on \(m\) inputs — one per gate — with \(\mathrm{depth}(C) \ge 1\), and consider the composite Boolean function \(x \mapsto C\bigl(g_0(x), \dots , g_{m-1}(x)\bigr)\). Suppose that each gate \(g_i\) admits both a DNF representation \(\varphi _i\) of width at most \(l\) computing the same function as \(g_i\) — one in which, within every term, any two literals sharing a variable coincide and no literal is repeated — and a CNF representation of width at most \(l\) computing that same function. Then there exist a layer-2 datum \(\mathcal{D}'\) on \(n\) variables of width at most \(l\), with gates \(g'_0, \dots , g'_{m'-1}\), and a Boolean circuit \(C'\) on \(m'\) inputs with \(\mathrm{depth}(C') \le \mathrm{depth}(C) - 1\), such that for every assignment \(x : \mathrm{Fin}\, n \to \mathbb {F}_2\),
Let \(\mathrm{data}\) be layer-2 data over \(n\) Boolean variables: a family of gates, each computing a DNF formula, sharing a common positive width bound. Let \(c_{\mathrm{top}}\) be a Boolean circuit whose input variables are indexed by these gates, and fix an integer \(l \ge 1\). Suppose that \(c_{\mathrm{top}}\) has depth exactly \(1\) (a single AND or OR gate applied to literals in the gate outputs), and that each gate computes a function that is representable both by a DNF of width at most \(l\) — one in which, within every term, distinct literals carry distinct variables and no literal is repeated — and by a CNF of width at most \(l\). Then there exist layer-2 data \(\mathrm{data}'\) of width at most \(l\) and a circuit \(c_{\mathrm{top}}'\) of depth \(0\) over the gates of \(\mathrm{data}'\) such that, for every Boolean assignment \(x\) to the \(n\) variables, the value of \(c_{\mathrm{top}}\) evaluated on the outputs of the gates of \(\mathrm{data}\) equals the value of \(c_{\mathrm{top}}'\) evaluated on the outputs of the gates of \(\mathrm{data}'\).
Let \(D\) be layer-2 data over \(n\) variables: a family of DNF gates \(g_0,\dots ,g_{m-1}\) together with a common positive width bound. Let \(C\) be a Boolean circuit whose \(m\) inputs are indexed by these gates and whose depth is at least \(2\), and let \(l\ge 1\). Suppose that for each index \(i\) the gate \(g_i\) admits a DNF formula of width at most \(l\) that computes the same Boolean function as \(g_i\) and in which, within every term, distinct literals use distinct variables and no literal is repeated, as well as a CNF formula of width at most \(l\) computing that same Boolean function. Then there exist layer-2 data \(D'\) over \(n\) variables of width at most \(l\), with gate family \(h_0,\dots ,h_{k-1}\), and a Boolean circuit \(C'\) whose \(k\) inputs are indexed by these gates and whose depth is at most \(\mathrm{depth}(C)-1\), such that for every assignment \(x\colon \mathrm{Fin}\, n\to \{ \mathrm{true},\mathrm{false}\} \) the value of \(C\) on the input vector \(\bigl(g_i(x)\bigr)_i\) equals the value of \(C'\) on the input vector \(\bigl(h_j(x)\bigr)_j\).
Let \(m\) be a natural number and let \(\ell _0,\dots ,\ell _{m-1}\) be natural numbers (the layer sizes) whose sum is at most a positive integer \(s\). Fix a real number \(\varepsilon \gt 0\), and suppose that \(b_0,\dots ,b_{m-1}\) are real numbers satisfying
Then \(\sum _{i} b_i \le \varepsilon /2\).
Let \(n\) and \(l\) be natural numbers, let \(\mathrm{gates}_0,\dots ,\mathrm{gates}_{s-1}\) be a family of DNF formulas in \(n\) variables, and let \(\rho \) be a restriction on the \(n\) variables. Suppose that for every index \(i\) the restricted function \(x \mapsto (\mathrm{gates}_i)(\rho .\mathrm{extend}\, x)\) has decision-tree depth at most \(l\). Then for every index \(i\) there is a CNF formula \(\psi \) of width at most \(l\) that computes this restricted function, that is, \(\psi (x) = (\mathrm{gates}_i)(\rho .\mathrm{extend}\, x)\) for every Boolean assignment \(x\).
Let \(cs\) be a finite list of Boolean circuits on \(n\) variables, let \(\rho \) be a restriction on those variables, and let \(l\) be a natural number. Suppose that for every circuit \(c\) in \(cs\), the restricted function \(x \mapsto c(\rho .\mathrm{extend}\, x)\) has decision-tree depth at most \(l\). Then there is a CNF formula \(\psi \) of width at most \(l\) that computes the restriction of the AND of \(cs\): writing \(g\) for the AND gate applied to the circuits in \(cs\), one has \(\psi (x) = g(\rho .\mathrm{extend}\, x)\) for every assignment \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\).
Fix natural numbers \(n\), \(s_2\), and \(l\), a restriction \(\rho _1\) on \(n\) variables, and a family of DNF formulas \(g_0, g_1, \dots , g_{s_2-1}\) on \(n\) variables. For each index \(i\), write \(\tilde g_i\) for the restricted function \(x \mapsto g_i(\rho _1.\mathrm{extend}\, x)\), obtained by evaluating \(g_i\) with the coordinates fixed by \(\rho _1\) set accordingly and the free coordinates supplied by \(x\). Suppose that every \(\tilde g_i\) has decision-tree depth at most \(l\). Then there is a CNF formula \(\Psi \) on \(n\) variables of width at most \(l\) in which each clause has no repeated literal and no two distinct literals sharing a variable, such that for every assignment \(x\),
Fix natural numbers \(m\), \(n\), and \(l\), and let \(g_0, \dots , g_{m-1}\) be Boolean functions of an input \(x : \mathrm{Fin}\, n \to \{ 0,1\} \), one for each index \(i \in \mathrm{Fin}\, m\), such that each \(g_i\) is computed both by some DNF of width at most \(l\) and by some CNF of width at most \(l\). Let the circuits \(c \in cs\) each be a single literal over the variables \(0, \dots , m-1\), and let \(C\) be the AND gate whose children are exactly these literals, evaluated on an input \(x\) by substituting the value \(g_i(x)\) for the \(i\)-th variable. Then the Boolean function computed in this way — the conjunction of the literal values \(g_i(x)\) or their negations — is itself represented by a CNF \(\psi \) of width at most \(l\), in the sense that \(\psi \) evaluates to \(C\)’s value on every input \(x\).
Fix \(n\) and a parameter \(p\) with \(0 \le p \le 1\), and let \(A\) be a predicate on restrictions of \(n\) variables. Writing \(\Pr [A]\) for the probability that \(A\) holds under a Bernoulli(\(p\)) random restriction, one has
where \(\neg A\) is the predicate that holds exactly for those restrictions on which \(A\) fails.
Let \(f, g : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be Boolean functions on \(n\) variables that agree at every input, and fix a real parameter \(p\) and a threshold \(t \in \mathbb {N}\). Then the Bernoulli(\(p\)) restriction probability of the event that a random restriction \(\rho \) produces a restricted function of decision-tree depth strictly greater than \(t\) is the same for \(f\) as for \(g\).
Fix a real parameter \(p\), a threshold \(t \in \mathbb {N}\), and two Boolean functions \(f, g : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) that agree at every input, so \(f(x) = g(x)\) for all \(x\). Draw a Bernoulli(\(p\)) random restriction \(\rho \) of the \(n\) variables, write \(h_\rho \) for the restricted function obtained from a Boolean function \(h\) under \(\rho \), and let \(D(\cdot )\) denote decision-tree depth. Then
i.e. the two functions assign the same probability to the event that the restricted function has decision-tree depth strictly greater than \(t\).
Let \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be a Boolean function, let \(t \in \mathbb {N}\) be a threshold, and let \(p_1, p_2 \in \mathbb {R}\) satisfy \(0 \lt p_1 \le 1\) and \(0 \lt p_2 \le 1\). For a parameter \(p\), let \(\Pr _p[\, \cdot \, ]\) denote the probability of an event under a Bernoulli(\(p\)) random restriction \(\rho \), in which each variable is left free with probability \(p\) and is otherwise fixed to a constant Boolean value, each of the two values with probability \((1-p)/2\). For a restriction \(\rho \), write \(f_\rho \) for the function of the free coordinates obtained from \(f\) by fixing each set coordinate \(i\) to the value \(\rho (i)\), and let \(D(f_\rho )\) be its decision-tree depth, that is, the least \(d\) for which some decision tree of depth at most \(d\) computes \(f_\rho \). Then
Fix a parameter \(p\) with \(0 \le p \le 1\), a Boolean function \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\), and thresholds \(l_1 \le l_2\). Under a Bernoulli(\(p\)) random restriction \(\rho \) on the \(n\) variables, the probability that the restricted function \(x \mapsto f(\rho .\mathrm{extend}\, x)\) has decision-tree depth strictly greater than \(l_2\) is at most the probability that it has decision-tree depth strictly greater than \(l_1\).
Fix a parameter \(p\) with \(0 \le p \le 1\), and consider a Bernoulli\((p)\) random restriction on \(n\) variables, under which each variable is independently left free with probability \(p\) and fixed to each of the two Boolean values with probability \((1-p)/2\). Let \(c_1, \dots , c_m\) be a list of Boolean circuits on \(n\) variables, and for each circuit \(c\) let \(\mathrm{bad}(c, \cdot )\) be a predicate on restrictions. Then the probability that some circuit in the list is bad is at most the sum of the individual probabilities:
where each term is the probability, under the same random restriction, that the corresponding predicate holds.
Fix a parameter \(p \in [0,1]\), and for a predicate \(E\) on restrictions of \(n\) variables write \(\mathbb {P}_p[E]\) for the probability that \(E\) holds under a Bernoulli(\(p\)) random restriction. If \(A\) and \(B\) are predicates on restrictions such that \(A(\rho )\) implies \(B(\rho )\) for every restriction \(\rho \), then
Fix a parameter \(p \in \mathbb {R}\) with \(0 \le p \le 1\), and let \(A_0, \dots , A_{m-1}\) be finitely many events on restrictions of \(n\) variables. Under a Bernoulli(\(p\)) random restriction, the probability that at least one of the events \(A_i\) occurs is at most the sum of the probabilities of the individual events:
where \(\mathbb {P}_p[\cdot ]\) denotes the Bernoulli event probability at parameter \(p\).
Let \(p\) be a real parameter and let \(\rho \) be a restriction on \(n\) variables, assigning each coordinate either a fixed Boolean value or leaving it free. Writing \(k\) for the number of free coordinates, the Bernoulli weight of \(\rho \) equals the product over all \(n\) coordinates of the per-variable weights:
where each factor \(w_p(\rho (i))\) equals \(p\) when coordinate \(i\) is free and \((1-p)/2\) when it is fixed.
Fix an arity \(k\), a gate type given by a flag that selects AND when true and OR when false, and a family of polarities \(\sigma : \mathrm{Fin}\, k \to \mathbb {F}_2\). For each variable index \(j \in \mathrm{Fin}\, k\), let \(\ell _j\) be the literal on variable \(j\) with polarity \(\sigma (j)\), and let \(G\) be the single gate of the chosen type applied to the list \(\ell _0, \ell _1, \dots , \ell _{k-1}\). Then there is a Boolean circuit \(c'\) of depth at most \(1\) that computes the same Boolean function as \(G\): for every assignment \(g : \mathrm{Fin}\, k \to \mathbb {F}_2\), the value of \(c'\) at \(g\) equals the value of \(G\) at \(g\).
Fix natural numbers \(m,n,l\). Let \(c_j\) be a Boolean circuit whose \(m\) inputs are indexed by \(\mathrm{Fin}\, m\) and whose depth is at most \(1\), and let \(\mathrm{gates}\) assign to each index \(i \in \mathrm{Fin}\, m\) a Boolean function \(\mathrm{gates}_i \colon (\mathrm{Fin}\, n \to \mathbb {F}_2) \to \mathbb {F}_2\). Suppose that every \(\mathrm{gates}_i\) is computed by some DNF over \(n\) variables of width at most \(l\), and also by some CNF over \(n\) variables of width at most \(l\). Then there exist a DNF \(\varphi \) over \(n\) variables of width at most \(l\) and a sign bit \(s \in \mathrm{Bool}\) such that, for every assignment \(x \colon \mathrm{Fin}\, n \to \mathbb {F}_2\), the value \(\varphi \) takes at \(x\) when \(s = \mathrm{true}\), and its negation when \(s = \mathrm{false}\), equals the value of \(c_j\) under the assignment \(i \mapsto \mathrm{gates}_i(x)\); moreover, each term of \(\varphi \) contains no repeated literal, and no two of its literals share a variable.
Let \(g\) be a Boolean circuit on \(n\) variables that is a single gate — an AND or OR gate applied to a list of child circuits — and let \(c\) be one of those children. If the size of \(g\) (its total number of gate and literal nodes) is at most \(s\), then the size of \(c\) is at most \(s - 1\).
The Boolean function on \(\mathrm{Fin}\, n \to \mathrm{Bool}\) computed by a subcircuit \(c_{\mathrm{sub}}\) when each of its inputs \(i\) is replaced by the gate function \(\mathrm{gates}\, i\), namely \(x \mapsto c_{\mathrm{sub}}.\mathrm{eval}\, (\lambda i,\ \mathrm{gates}\, i\, x)\).
Let \(C\) be a Boolean circuit on \(n\) variables of the form of a single gate — an AND gate or an OR gate — applied to a list of child circuits, and let \(c\) be one of those children. If the depth of \(C\) is at most \(d\), then the depth of \(c\) is at most \(d - 1\).
Let \(g\) be an internal gate of a Boolean circuit on \(n\) variables — either an AND gate or an OR gate — applied to a list of child circuits, and suppose its maximum fanin is at most \(w\). Then every child circuit occurring in that list also has maximum fanin at most \(w\).
Let \(g = \mathrm{node}\, (\mathit{isAnd},\, cs)\) be an internal gate of a Boolean circuit on \(n\) variables, with \(cs\) the list of its child subcircuits, and let \(s\) be a natural number with \(\mathrm{size}(g) \le s\). Then the sum of the sizes of the children, \(\sum _{c \in cs} \mathrm{size}(c)\), is at most \(s - 1\) (truncated natural-number subtraction).
Let \(c\) be a Boolean circuit on \(m\) variables — a tree whose leaves are literals and whose internal nodes are AND or OR gates applied to lists of subcircuits — and suppose its depth is \(0\). Then \(c\) is a single leaf: there exists a literal \(l\) such that \(c\) consists of that literal alone.
Let \(f\colon \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function on \(n \ge 1\) variables, and let \(d,s,w,l,t \in \mathbb {N}\) satisfy \(d \ge 2\) and \(s,w,l \ge 1\). Suppose \(f\) is computed by some Boolean circuit of depth at most \(d\), size at most \(s\), and maximum fan-in at most \(w\). Write \(f\! \restriction _\rho \) for the restriction of \(f\) by a restriction \(\rho \), and let \(D(g)\) denote the decision-tree depth of a Boolean function \(g\), that is, the least depth of a decision tree computing it. Then, under a Bernoulli random restriction \(\rho \) with parameter \(\delta = \dfrac {1}{40w}\left(\dfrac {1}{40l}\right)^{d-2}\),
Let \(f:\{ 0,1\} ^n\to \{ 0,1\} \) be a Boolean function on \(n\ge 1\) variables, and let \(d,s,w,l,t\) be natural numbers with \(d\ge 2\) and \(s,w,l\ge 1\). Suppose \(f\) is computed by some Boolean circuit of depth at most \(d\), size at most \(s\), and fan-in at most \(w\). Draw a random restriction \(\rho \) on the \(n\) variables in which each variable is independently left free with probability
and is otherwise fixed to \(0\) or to \(1\), each with probability \((1-p)/2\); let \(f\restriction _\rho \) denote the function of the free variables obtained by substituting the fixed values into \(f\). Then the probability that \(f\restriction _\rho \) cannot be computed by any decision tree of depth at most \(t\) satisfies
Fix \(n \ge 1\), and let \(f\colon \{ 0,1\} ^n \to \{ 0,1\} \) be the conjunction of \(s_2\) DNF formulas \(g_1,\dots ,g_{s_2}\), so that \(f(x)=1\) if and only if \(g_i(x)=1\) for every \(i\). Assume a width bound \(w \ge 1\) under which each \(g_i\) has width at most \(w\) (every term of every gate has at most \(w\) literals), and assume that within each term of each gate the literals are distinct and no variable occurs twice. Fix integers \(l \ge 1\) and \(t\), and let \(\rho \) be a Bernoulli random restriction with parameter \(\frac{1}{40w}\cdot \frac{1}{40l}\); write \(p_1=\frac{1}{40w}\) and \(p_2=\frac{1}{40l}\). Write \(f\! \restriction _\rho \) for the function obtained by extending each input with the values \(\rho \) fixes, and let \(D(\cdot )\) denote decision-tree depth (the least depth of a decision tree computing the given function). Then the probability that the restricted function has decision-tree depth exceeding \(t\) obeys
Let \(f:\{ 0,1\} ^n\to \{ 0,1\} \) with \(n\ge 1\), and suppose \(f\) is the conjunction of \(s_2\) DNF formulas \(D_1,\dots ,D_{s_2}\), so that \(f(x)=1\) exactly when \(D_i(x)=1\) for every index \(i\); equivalently, \(f\) is computed by a depth-three AND-of-OR-of-AND circuit. Assume each formula \(D_i\) has width at most \(w\) with \(w\ge 1\), and that within every term no literal is repeated and no variable occurs in two opposite polarities, so that a term reads each variable at most once. Fix integers \(l\ge 1\) and \(t\ge 0\) and a real number \(\varepsilon \) satisfying \(s_2\, (1/2)^{l}\le \varepsilon /2\) and \((1/2)^{t}\le \varepsilon /2\), and put \(p_1=1/(40w)\) and \(p_2=1/(40l)\). Draw a Bernoulli random restriction \(\rho \) with parameter \(p_1 p_2\), and write \(f_\rho \) for the restriction of \(f\) by \(\rho \). Then the probability that the decision-tree depth of \(f_\rho \) exceeds \(t\) is at most
Let \(f\colon \{ 0,1\} ^n\to \{ 0,1\} \) be a Boolean function on \(n\ge 1\) variables, and let \(d\ge 2\), \(s\ge 1\), \(w\ge 1\), and \(l\ge 1\) be positive integers together with an integer threshold \(t\ge 0\). Suppose \(f\) is computed by some Boolean circuit of depth at most \(d\), size at most \(s\) (counting gates and literal leaves), and fan-in at most \(w\). Let \(\rho \) be a Bernoulli random restriction on the \(n\) variables with survival parameter
so that each variable is independently left free with probability \(p\) and otherwise fixed to \(0\) or to \(1\), each with probability \((1-p)/2\). Write \(f\restriction _\rho \) for the function obtained from \(f\) by substituting the values \(\rho \) assigns and keeping the free coordinates as inputs, and let \(D(g)\) denote the least depth of a decision tree computing a Boolean function \(g\). Then
Let \(f\colon \{ 0,1\} ^n\to \{ 0,1\} \) be a Boolean function on \(n\ge 1\) variables, let \(s,w,l\ge 1\) be positive integers, let \(t\ge 0\), and suppose \(f\) is computed by a Boolean circuit of depth at most \(2\), size at most \(s\), and fan-in at most \(w\). Draw a random restriction \(\rho \) that leaves each variable free independently with probability \(p=\tfrac {1}{40w}\) and otherwise fixes it to \(0\) or \(1\), each with probability \(\tfrac {1-p}{2}\). Writing \(D(g)\) for the least depth of a decision tree computing a Boolean function \(g\), and \(f\! \restriction _\rho \) for the function obtained from \(f\) by substituting the fixed values of \(\rho \) and keeping the free coordinates as inputs, the probability that the restricted function is hard to compute by a shallow decision tree satisfies
Let \(f:\{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function on \(n \ge 1\) variables, and fix natural numbers \(d,s,w,l,t\) with \(d \ge 3\) and \(s,w,l \ge 1\). Suppose \(f\) is computed by a Boolean circuit of depth at most \(d\), size at most \(s\), and maximum fan-in at most \(w\). For parameters \(w,l\) and a depth \(e\), write \(\delta _e = \frac{1}{40w}\left(\frac{1}{40l}\right)^{e-2}\) for the composed restriction parameter; let \(D(g)\) denote the decision-tree depth of a Boolean function \(g\), and let \(g|_\rho \) denote its restriction by a restriction \(\rho \). Assume the following inductive hypothesis: for every Boolean function \(g\) on \(n\) variables and every integer \(s' \ge 1\), if \(g\) is computed by a circuit of depth at most \(d-1\), size at most \(s'\), and maximum fan-in at most \(w\), then under a Bernoulli random restriction \(\rho \) with parameter \(\delta _{d-1}\),
Then under a Bernoulli random restriction \(\rho \) with parameter \(\delta _{d}\),
Let \(c\) be a clause on the variables \(x_0,\dots ,x_{n-1}\), that is, a list of literals read as their disjunction, where a literal is a variable together with a sign and evaluates to that variable’s value if positive and to its negation if negated. Suppose \(c\) is tautological: it contains two literals on the same variable with opposite signs. Then for every Boolean assignment \(x \colon \{ 0,\dots ,n-1\} \to \{ \text{true},\text{false}\} \), at least one literal of \(c\) evaluates to \(\mathrm{true}\), so the disjunction of the literals of \(c\) evaluates to \(\mathrm{true}\).
Let \(\psi \) be a CNF formula over \(n\) Boolean variables — a conjunction of clauses, each clause a disjunction of literals — and let \(x \in \{ 0,1\} ^n\) be an assignment. Form the cleaned formula from \(\psi \) by first deleting every tautological clause (a clause containing two literals on the same variable with opposite signs) and then, within each surviving clause, keeping only the first literal on each variable and discarding later literals on variables already seen. Then the cleaned formula and \(\psi \) evaluate to the same Boolean value on \(x\).
Let \(\psi \) be a CNF formula over \(n\) variables, and let \(\mathrm{clean}(\psi )\) be the formula obtained from \(\psi \) by discarding every tautological clause and then, in each surviving clause, keeping only the first literal on each variable. Then every clause of \(\mathrm{clean}(\psi )\), viewed as a list of literals, contains no repeated literal.
Let \(\psi \) be a CNF formula over the variables \(\mathrm{Fin}\, n\). After cleaning \(\psi \) — that is, after discarding every tautological clause and then deduplicating the variables within each surviving clause — every clause becomes variable-injective: in any clause of the cleaned formula, any two literals sharing the same variable are in fact the same literal.
Let \(\psi \) be a CNF formula over \(n\) variables, and form its cleaned version by discarding every tautological clause (one containing two literals on the same variable with opposite signs) and then, within each surviving clause, keeping only the first occurrence of each variable. Writing the width of a CNF formula for the largest number of literals in any of its clauses (and \(0\) for the empty formula), the width of the cleaned formula is at most the width of \(\psi \).
Represent a CNF formula over the variables indexed by \(\{ 0,1,\dots ,n-1\} \) as a conjunction of clauses, each clause being a disjunction of literals, where a literal is a variable or its negation. Cleaning such a formula proceeds in two steps: first delete every clause that contains some variable together with both polarities (one literal negated, one not), and then, within each surviving clause, keep a literal only when its variable has not already occurred among the literals retained earlier in that clause. Then for every CNF formula \(c\) and every Boolean assignment \(x \colon \{ 0,\dots ,n-1\} \to \{ \mathrm{true},\mathrm{false}\} \), the cleaned formula and \(c\) take the same Boolean value at \(x\).
Let \(c\) be a CNF formula over \(n\) variables, and form its cleaned version by first discarding every clause that contains some variable both positively and negatively, and then deleting from each surviving clause every literal whose variable already occurs earlier in that clause. Then each clause of the resulting formula, viewed as a list of literals, contains no repeated literal.
Let \(c\) be a CNF formula over the variable set \(\{ 0,\dots ,n-1\} \), and let \(\hat c\) be the formula obtained from \(c\) by first deleting every clause that contains some variable together with its negation, and then, in each surviving clause, keeping only one literal per variable. Then within any single clause of \(\hat c\) the variable of a literal determines that literal completely: if two literals of the clause share the same variable, they are equal (and in particular have the same polarity).
Let \(c\) be a CNF formula, that is, a list of clauses, each clause being a list of literals over the variables \(\{ 0,\dots ,n-1\} \). Define the width of a clause to be its number of literals and the width of \(c\) to be the maximum width among its clauses (and \(0\) if \(c\) has no clauses). Form the cleaned formula from \(c\) by discarding every clause that contains some variable together with both of its polarities, and then, within each surviving clause, keeping only the first occurrence of each variable and deleting all later literals on a variable already seen. Then the width of the cleaned formula is at most the width of \(c\).
Let \(d\) be a DNF formula in \(n\) Boolean variables — a disjunction of terms, each term a conjunction of literals over the variables \(\{ 0,\dots ,n-1\} \) — and form the cleaned formula by first discarding every term that contains some variable together with its negation, and then, in each surviving term, keeping only the first literal on each variable and deleting any later literal repeating that variable. Then the cleaned formula and \(d\) compute the same Boolean function: for every assignment \(x \colon \{ 0,\dots ,n-1\} \to \{ \mathrm{true},\mathrm{false}\} \), the cleaned formula evaluates to true at \(x\) if and only if \(d\) does.
Let \(d\) be a DNF formula in \(n\) variables, and clean it by first deleting every term that contains some variable together with its negation, and then, within each surviving term, keeping only the first literal on each variable and discarding any later literal whose variable has already appeared. Every term of the resulting formula is then duplicate-free: no literal occurs twice in it, where two literals count as equal when they share both variable and polarity.
Let \(d\) be a DNF formula over \(n\) variables — a list of terms, each term a list of literals, where a literal names a variable in \(\{ 0,\dots ,n-1\} \) together with a polarity (negated or not). Form a new formula \(d'\) from \(d\) by first discarding every term that contains two literals on the same variable with opposite polarities, and then, in each surviving term, keeping only the first literal encountered on each variable and deleting all later literals that repeat an already-kept variable. Then in every term of \(d'\) the variable determines the literal: for any two literals \(l_1, l_2\) occurring in such a term, if \(l_1\) and \(l_2\) name the same variable, then \(l_1 = l_2\).
Let \(d\) be a DNF formula over \(n\) variables — a disjunction of terms, each term a conjunction of literals — and form its cleaned version by first deleting every contradictory term (one containing a variable together with its negation) and then, in each surviving term, discarding any literal whose variable already occurs in an earlier retained literal of that term. The width of a term is its number of literals, and the width of a formula is the largest width of any of its terms (with width \(0\) for the empty formula). Then the width of the cleaned formula is at most the width of \(d\).
Let \(\psi _1, \dots , \psi _k\) be CNF formulas over the variables \(\{ 1, \dots , n\} \), and form the single CNF formula obtained by concatenating all of their clause lists into one. Then for every assignment \(x \colon \{ 1,\dots ,n\} \to \{ \text{true},\text{false}\} \), this concatenated formula evaluates to true at \(x\) if and only if every \(\psi _i\) evaluates to true at \(x\).
Let \(\psi _1, \dots , \psi _s\) be CNF formulas over \(n\) variables, and let \(\ell \) be a natural number such that each \(\psi _i\) has width at most \(\ell \). Then the CNF formula obtained by concatenating all of their clause lists into a single list has width at most \(\ell \).
Fix parameters \(p, q \in \mathbb {R}\) and a restriction \(\sigma \) on \(n\) variables. Summing the product of the Bernoulli weight of \(\rho _1\) at parameter \(p\) and the Bernoulli weight of \(\rho _2\) at parameter \(q\) over all pairs of restrictions \((\rho _1, \rho _2)\) on \(n\) variables whose composition equals \(\sigma \) yields the Bernoulli weight of \(\sigma \) at parameter \(pq\):
where \(w_r(\rho )\) denotes the Bernoulli weight of \(\rho \) at parameter \(r\), \(\rho _1 \circ \rho _2\) is the restriction that fixes each coordinate to the value given by \(\rho _1\) where defined and to that of \(\rho _2\) otherwise, and \([\, \cdot \, ]\) is the indicator taking the value \(1\) when its condition holds and \(0\) otherwise.
Let \(w \ge 1\) be a natural number, let \(l \ge 1\) be a real number, and let \(d \ge 2\) be a natural number. Then the composed Bernoulli restriction parameter satisfies
Let \(w\) be a positive integer fan-in bound, let \(l \gt 0\) be a width parameter, and let \(d\) be any depth. Then the composed Bernoulli restriction parameter
is strictly positive.
Write \(\delta (w,l,d) = \dfrac {1}{40w}\left(\dfrac {1}{40l}\right)^{d-2}\) for the composed restriction parameter, with fan-in bound \(w\), width parameter \(l\), and depth \(d\). For natural numbers \(w\) and \(l\) and any depth \(d \ge 3\),
Fix a fan-in bound \(w\), a width parameter \(l \gt 0\), and a depth \(d \ge 3\), and for these values write \(\delta (w,l,d)\) for the composed Bernoulli restriction parameter, namely \(\delta (w,l,d) = \frac{1}{40w}\left(\frac{1}{40l}\right)^{d-2}\). Then decreasing the depth by one removes a single factor of \(1/(40l)\):
Given two restrictions \(\rho _1, \rho _2\) on \(n\) variables, their composition is the restriction that on each coordinate \(i\) uses the value fixed by \(\rho _1\) if there is one, and otherwise the value given by \(\rho _2\). This models applying \(\rho _1\) first and then \(\rho _2\) on the variables left free by \(\rho _1\).
Let \(\rho _1,\rho _2,\sigma \) be restrictions on \(n\) variables, and let \(\rho _1 \ast \rho _2\) denote the composition of \(\rho _1\) and \(\rho _2\), whose value at a coordinate \(i\) is the value fixed by \(\rho _1\) at \(i\) if \(\rho _1\) fixes it and otherwise the value \(\rho _2\) assigns at \(i\). Then \(\rho _1 \ast \rho _2 = \sigma \) if and only if, for every coordinate \(i \in \mathrm{Fin}\, n\), the value of \(\rho _1 \ast \rho _2\) at \(i\) equals \(\sigma (i)\).
Let \(\rho \) be a restriction on \(n\) variables, and let \(\varnothing \) denote the empty restriction that leaves every variable free. Then the composition of \(\rho \) with \(\varnothing \), applying \(\rho \) first and then \(\varnothing \) on the variables \(\rho \) leaves free, equals \(\rho \).
Fix \(n \ge 1\), an integer \(l \ge 1\), and a natural number \(t\), and write \(D(g)\) for the decision-tree depth of a Boolean function \(g\) on \(n\) variables, that is, the least depth of a decision tree computing \(g\). For a function \(g\) and a restriction \(\rho \) on the \(n\) variables, let \(g\! \restriction _{\rho }\) denote the function obtained from \(g\) by fixing each coordinate to the value \(\rho \) assigns it and leaving the coordinates \(\rho \) marks free as inputs. Let \(c_1,\dots ,c_m\) be Boolean circuits on \(n\) variables, and let \(F\) be the function computed by a single AND gate, or by a single OR gate, applied to \(c_1,\dots ,c_m\); the bound below holds for either gate. Suppose a restriction \(\rho _1\) satisfies \(D\! \left(c_i\! \restriction _{\rho _1}\right) \le l\) for every \(i\). Draw a random restriction \(\rho _2\) under which each of the \(n\) variables is independently left free with probability \(p = \tfrac {1}{40 l}\) and otherwise fixed, taking each of the values \(0\) and \(1\) with probability \(\tfrac {1-p}{2}\), and let \(\rho _1 \circ \rho _2\) be the composite restriction that keeps every value fixed by \(\rho _1\) and applies \(\rho _2\) to the coordinates \(\rho _1\) leaves free. Then
Fix \(n\) and a width bound \(\ell \), and let \(f_1, \dots , f_k\) be Boolean functions on \(n\) variables, each mapping an assignment \(x \in \{ 0,1\} ^n\) to a truth value. Suppose that every \(f_i\) is computed by some CNF formula of width at most \(\ell \), that is, a conjunction of clauses in which each clause has at most \(\ell \) literals. Then their conjunction \(x \mapsto \bigwedge _{i=1}^{k} f_i(x)\) is also computed by a single CNF formula of width at most \(\ell \).
Let \(f_1, \dots , f_m\) be Boolean functions on \(\{ 0,1\} ^n\), and let \(\ell \) be a natural number. Suppose that each \(f_i\) is computed by some DNF formula of width at most \(\ell \), that is, a disjunction of terms in which every term has at most \(\ell \) literals. Then their disjunction \(x \mapsto f_1(x) \lor \cdots \lor f_m(x)\) is also computed by a single DNF formula of width at most \(\ell \).
Let \(t\) be a term over \(n\) Boolean variables — a list of literals, each a variable together with a polarity — and suppose \(t\) contains two literals on the same variable with opposite polarities, one negated and one not. Then, read as a single clause under disjunction semantics, \(t\) evaluates to true under every assignment \(x : \{ 0,\dots ,n-1\} \to \{ \mathrm{true},\mathrm{false}\} \); that is, some literal of \(t\) holds under \(x\), for every \(x\).
Let \(t\) be a term (a conjunction of literals) over the variables \(x_0,\dots ,x_{n-1}\). If \(t\) contains two literals on the same variable with opposite polarities — one negated and one not — then \(t\) evaluates to false under every assignment \(x : \{ 0,\dots ,n-1\} \to \{ \mathrm{true},\mathrm{false}\} \).
Let \(c\) be a clause over the variables \(x_0,\dots ,x_{n-1}\) — a list of literals read as their disjunction, where a literal evaluates to \(x_i\) if positive and \(\lnot x_i\) if negated. Suppose \(c\) is not tautological, meaning it contains no two literals on the same variable with opposite signs. Then for every Boolean assignment \(x \colon \{ 0,\dots ,n-1\} \to \{ \mathrm{false},\mathrm{true}\} \), the clause obtained from \(c\) by keeping only the first literal on each variable evaluates to the same truth value as \(c\) itself.
Let \(c\) be a clause, that is, a finite list of literals over the variables \(\{ 0,1,\dots ,n-1\} \), and let \(\mathrm{dedup}(c)\) be the clause obtained from \(c\) by retaining, for each variable, only its first occurring literal and discarding every later literal on the same variable. Then the length of \(\mathrm{dedup}(c)\) is at most the length of \(c\).
Let \(c\) be a finite list of literals over the variable set \(\{ 0,1,\dots ,n-1\} \), and form a new list from \(c\) by scanning left to right and discarding every literal whose variable has already appeared, keeping only the first occurrence of each variable. Then this deduplicated list is repetition-free: no literal occurs in it more than once.
Let \(c\) be a clause, that is, a list of literals over the variables \(\{ 0,1,\dots ,n-1\} \), each literal consisting of a variable together with a negation flag. Form the deduplicated clause obtained from \(c\) by keeping, for each variable, only its first occurrence. Then any two literals of this deduplicated clause that share the same variable are in fact the same literal; equivalently, distinct literals of the deduplicated clause have distinct variables.
Let \(t\) be a term over \(n\) variables, that is, a list of literals, each a variable in \(\{ 0,\dots ,n-1\} \) together with a negation flag. Form a new term \(t'\) by scanning \(t\) and retaining a literal exactly when no already-retained literal has the same variable, discarding any literal whose variable has already occurred. Then \(t'\) has no repeated literal: no literal occurs twice in \(t'\).
Let \(t\) be a term over the variables \(x_0,\dots ,x_{n-1}\), that is, a list of literals, and let \(x \in \{ 0,1\} ^n\) be an assignment. Suppose \(t\) contains no two literals on the same variable with opposite polarities. Let \(t'\) be the term obtained from \(t\) by scanning it and keeping a literal only when no earlier-kept literal uses the same variable, so that each variable of \(t\) occurs at most once in \(t'\). Then \(t'\) and \(t\) have the same value as a clause under \(x\): some literal of \(t'\) is satisfied by \(x\) if and only if some literal of \(t\) is.
Let \(t\) be a term over \(n\) variables, that is, a list of literals \(l\), each carrying a variable \(l.\mathrm{var} \in \{ 0,\dots ,n-1\} \) and a polarity flag, and interpreted conjunctively: \(t\) evaluates to true on an assignment \(x \in \{ 0,1\} ^n\) exactly when every literal in \(t\) holds under \(x\). Suppose \(t\) is consistent, meaning it contains no two literals on the same variable with opposite polarities. Let \(t'\) be the term obtained from \(t\) by keeping, for each variable, only its first occurrence and discarding every later literal whose variable has already appeared. Then \(t'\) and \(t\) have the same value on every assignment \(x\).
Let \(t\) be a term over \(n\) variables, that is, a list of literals, each literal consisting of a variable in \(\{ 0,1,\dots ,n-1\} \) together with a negation flag. Let \(t'\) be the de-duplicated term obtained by scanning \(t\) and retaining each literal only when no earlier-retained literal already uses the same variable. Then within \(t'\) the variable of a literal determines that literal: any two literals \(l_1, l_2 \in t'\) whose variables agree are equal, that is, they have the same variable and the same negation flag.
Let \(t\) be a term, that is, a finite list of literals over the variables \(\{ 0, 1, \dots , n-1\} \), and let \(t'\) be the term obtained from \(t\) by scanning it and keeping each literal only when no literal already kept uses the same variable. Then \(t'\) has length at most that of \(t\).
Let \(f\colon \{ 0,1\} ^n\to \{ 0,1\} \) be a Boolean function on \(n\ge 1\) variables that is computed by some Boolean circuit of depth at most \(2\), size at most \(s\), and fan-in at most \(w\), where \(w\ge 1\). Fix a real parameter \(p\) with \(0\lt p\le \tfrac {1}{40w}\) (so in particular \(p\le 1\)), and draw a random restriction \(\rho \) in which each of the \(n\) variables is, independently, left free with probability \(p\) and otherwise fixed to the value \(0\) or to the value \(1\), each with probability \(\tfrac {1-p}{2}\). Write \(f_\rho \) for the restriction of \(f\) by \(\rho \) (the function \(x\mapsto f\) evaluated at the assignment agreeing with \(\rho \) on its fixed coordinates and with \(x\) on the free ones), and let \(D(g)\) denote the least depth of a decision tree computing a function \(g\). Then for every \(t\in \mathbb {N}\),
Converts a list of children \(cs\) of a top AND gate (of depth \(\le 2\)) into a CNF: each literal child becomes a singleton clause, each OR child becomes the clause of its literals, and each AND child contributes the singleton clauses of its literals.
Let \(cs\) be a list of Boolean circuits on \(n\) variables, viewed as the children of a top-level AND gate, and suppose the AND-of-children circuit \(\bigwedge cs\) has depth at most \(2\). Then the depth-\(2\) AND-to-CNF conversion applied to \(cs\) produces a CNF formula computing the same Boolean function as \(\bigwedge cs\): for every assignment \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\), the value of the resulting CNF at \(x\) equals the value of \(\bigwedge cs\) at \(x\).
Let \(cs\) be a list of Boolean circuits on \(n\) variables, regarded as the children of a top AND gate, and suppose this AND gate has depth at most \(2\). Then the CNF formula obtained by converting these children — each literal child becoming a singleton clause, each OR child the clause of its literals, and each AND child contributing the singleton clauses of its literals — has width at most the maximum fanin of the AND gate.
Converts a list of children \(cs\) of a top OR gate (of depth \(\le 2\)) into a DNF: each literal child becomes a singleton term, each AND child becomes the term of its literals, and each OR child contributes the singleton terms of its literals.
Let \(cs\) be a finite list of Boolean circuits on \(n\) variables, regarded as the children of a top-level OR gate \(\varphi \), and suppose \(\varphi \) has depth at most \(2\). Form a DNF \(D\) from \(cs\) by turning each literal child into a singleton term, each AND child into the single term consisting of all its literals, and each OR child into the collection of singleton terms of its literals (in each case reading a circuit literal as the corresponding DNF literal, keeping the same variable and polarity). Then \(D\) and \(\varphi \) compute the same Boolean function: \(D(x) = \varphi (x)\) for every assignment \(x \colon \{ 0,\dots ,n-1\} \to \{ \mathrm{true},\mathrm{false}\} \).
Let \(cs\) be a list of Boolean circuits on \(n\) variables, regarded as the children of a top-level OR gate, and suppose that this OR gate has depth at most \(2\). Form the DNF obtained from \(cs\) by converting each literal child into a singleton term, each AND child into the term consisting of its literals, and each OR child into the singleton terms of its literals. Then the width of this DNF (the maximum number of literals in any of its terms) is at most the maximum fan-in of the OR gate.
Fix \(n \ge 1\), and let \(D_1,\dots ,D_{s_2}\) be DNF formulas on \(n\) variables, each of width at most \(w\) with \(w \ge 1\), subject to the two syntactic conditions that within every term any two literals on the same variable coincide (so no term repeats a variable with opposite polarities) and that no term lists the same literal twice. Let \(l \in \mathbb {N}\), and let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le 1/(40w)\) and \(p \le 1\). Draw a Bernoulli\((p)\) random restriction \(\rho \) on the \(n\) variables. Then with probability at least
there exists a CNF formula \(\Psi \) of width at most \(l\) that computes the conjunction of the restricted gates: writing \(D_i\! \restriction _\rho \) for the function \(x \mapsto D_i(\rho .\mathrm{extend}\, x)\) obtained by fixing the coordinates set by \(\rho \) and leaving the free coordinates to \(x\), one has \(\Psi (x) = \bigwedge _{i=1}^{s_2} \bigl(D_i\! \restriction _\rho \bigr)(x)\) for every assignment \(x\).
Let \(g_1,\dots ,g_s\) be DNF formulas in \(n\) Boolean variables, and let \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be their conjunction, so that \(f(x) = \mathrm{true}\) if and only if \(g_i(x) = \mathrm{true}\) for every \(i\). Fix a restriction \(\rho \) of the \(n\) variables, and write \(h_\rho \) for the restriction of a Boolean function \(h\) by \(\rho \) (the function obtained by fixing each coordinate that \(\rho \) sets and leaving the rest free). Suppose that for every \(i\) the restricted gate \((g_i)_\rho \) has decision-tree depth at most \(l\). Then the restricted function \(f_\rho \) is computed by some CNF formula \(\Psi \) of width at most \(l\) — every clause contains at most \(l\) literals — in which each clause is simple: no literal occurs twice, and any two literals in a clause that share a variable are identical. That is, \(\Psi (x) = f_\rho (x)\) for all \(x\).
Let \(f:\{ 0,1\} ^n\to \{ 0,1\} \) be a Boolean function on \(n\ge 1\) variables, and write \(D(h)\) for the decision-tree depth of a Boolean function \(h\), i.e. the least depth of a decision tree computing it. Suppose \(f\) is the conjunction \(f=g_1\wedge \cdots \wedge g_{s_2}\) of DNF formulas \(g_1,\dots ,g_{s_2}\), meaning \(f(x)=1\) exactly when \(g_i(x)=1\) for every \(i\). Fix a restriction \(\rho _1\) on the \(n\) variables such that each restricted gate satisfies \(D(g_i\! \restriction \! \rho _1)\le l\), where \(l\ge 1\), and let \(t\) be a natural number. Suppose \(0\lt p_2\le \tfrac {1}{40l}\) (so in particular \(p_2\le 1\)), and let \(\rho _2\) be a Bernoulli(\(p_2\)) random restriction — each variable independently left free with probability \(p_2\) and fixed to \(0\) or \(1\) with probability \((1-p_2)/2\) each. Writing \(\rho _1\circ \rho _2\) for the composite restriction that applies \(\rho _1\) first and then \(\rho _2\) on the coordinates left free by \(\rho _1\), we have
Let \(n \ge 1\), and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be the conjunction of \(s_2\) DNF formulas \(g_1, \dots , g_{s_2}\) on \(n\) variables, so that \(f(x) = \mathrm{true}\) exactly when \(g_i(x) = \mathrm{true}\) for every \(i\). Suppose each \(g_i\) has width at most \(w\), where \(w \ge 1\), and that within every term of every \(g_i\) no literal is repeated and no two distinct literals share the same variable. Fix natural numbers \(l \ge 1\) and \(t\), together with real parameters \(p_1, p_2 \in (0,1]\) satisfying \(p_1 \le 1/(40w)\) and \(p_2 \le 1/(40l)\). Writing \(D(g)\) for the decision-tree depth of a Boolean function \(g\) and \(f\restriction _\rho \) for the restricted function \(x \mapsto f(\rho \text{-extension of }x)\), a Bernoulli\((p_1 p_2)\) random restriction \(\rho \) then satisfies
Let \(C\) be a Boolean circuit on \(n\) variables of the form \(g(c_1,\dots ,c_k)\), a single gate \(g\) (either AND or OR) applied to a list of child circuits \(c_1,\dots ,c_k\). If the depth of \(C\) is at most \(1\), then every child \(c_i\) is a literal, that is, \(c_i = \ell _i\) for some literal \(\ell _i\) (a variable index together with a polarity).
Let \(g = \mathrm{node}(\mathit{isAnd}, cs)\) be a Boolean circuit consisting of a single gate (an AND gate if \(\mathit{isAnd}\) is true, an OR gate otherwise) applied to a list \(cs\) of child circuits, and suppose its depth satisfies \(\mathrm{depth}(g) \le 2\). Then every child circuit \(c\) in \(cs\) has depth \(\mathrm{depth}(c) \le 1\).
Let \(\varphi _1, \dots , \varphi _k\) be DNF formulas over \(n\) variables, and let \(x \in \{ 0,1\} ^n\) be a Boolean assignment. Form the single DNF whose term-list is the concatenation of the term-lists of \(\varphi _1, \dots , \varphi _k\). Then this concatenated formula evaluates to true at \(x\) if and only if at least one of the formulas \(\varphi _i\) evaluates to true at \(x\).
Let \(\varphi _1, \dots , \varphi _s\) be DNF formulas over \(n\) variables, and let \(\ell \) be a natural number such that each \(\varphi _i\) has width at most \(\ell \) (that is, every term of every \(\varphi _i\) contains at most \(\ell \) literals). Then the DNF formula obtained by concatenating all of their term lists into a single list of terms also has width at most \(\ell \).
Let \(f\) be a Boolean function of \(n\) variables, and let \(\rho _1\) and \(\rho _2\) be two restrictions on those variables, each pinning some coordinates to a fixed Boolean value and leaving the rest free. Let \(\rho _1 \circ \rho _2\) be the composed restriction that fixes each coordinate \(i\) to the value assigned by \(\rho _1\) when \(\rho _1\) pins \(i\), and otherwise to the value assigned by \(\rho _2\). For a restriction \(\rho \), write \(f|_{\rho }\) for the function obtained from \(f\) by substituting the fixed values of \(\rho \) and treating the coordinates left free by \(\rho \) as its inputs, and let \(\operatorname {depth}_{\mathrm{DT}}(g)\) denote the least \(d\) for which some decision tree of depth at most \(d\) computes \(g\). Then further restricting along \(\rho _2\) can only lower this depth:
Let \(f\) and \(g\) be Boolean functions of \(n\) variables, that is, maps \(\{ 0,1\} ^n \to \{ 0,1\} \), and suppose they agree on every input, so \(f(x) = g(x)\) for all \(x \in \{ 0,1\} ^n\). Then \(f\) and \(g\) have the same decision-tree depth.
Let \(f\colon (\mathrm{Fin}\, n \to \mathbb {F}_2) \to \{ 0,1\} \) be a Boolean function on \(n\) variables, and let \(d\) be a natural number. If the decision-tree depth of \(f\) is at most \(d\), then there is a CNF formula \(\psi \) that computes \(f\) (that is, \(\psi (x) = f(x)\) for every assignment \(x\)) and whose width is at most \(d\), meaning every clause has at most \(d\) literals. Moreover \(\psi \) can be taken clean: within each clause no literal is repeated, and no two literals share the same underlying variable.
Let \(f : (\mathrm{Fin}\, n \to \mathbb {F}_2)\) — more precisely \(f\) a Boolean function on \(n\) variables — have decision-tree depth at most \(d\). Then \(f\) is computed by a DNF formula \(\varphi \) (a disjunction of terms, each term a conjunction of literals) of width at most \(d\), meaning every term of \(\varphi \) contains at most \(d\) literals, such that \(\varphi (x) = f(x)\) for all inputs \(x\). Moreover each term of \(\varphi \) is tidy: no literal is repeated within a term, and no two literals in the same term share a variable (equivalently, within any one term each variable occurs at most once, and with a single fixed polarity).
Let \(f\) be a Boolean function of \(n\) Boolean variables, and let \(\rho \) be a restriction that fixes some of these variables to Boolean values and leaves the remaining ones free. Write \(f_\rho \) for the restricted function, whose value on an input \(x\) is \(f\) evaluated at the assignment that agrees with \(\rho \) on the fixed coordinates and with \(x\) on the free coordinates. Writing \(D(g)\) for the decision-tree depth of a Boolean function \(g\) — the least depth of a decision tree computing \(g\) — one has \(D(f_\rho ) \le D(f)\).
Given a decision tree \(T\) and a partial assignment \(\rho \), the restricted tree \(\mathrm{dtRestrict}\, T\, \rho \) is obtained by collapsing every branch whose query variable is fixed by \(\rho \) into the corresponding subtree, and keeping unrestricted branches intact.
Let \(T\) be a decision tree on \(n\) Boolean variables, and let \(\rho \) be a restriction on those variables, that is, an assignment giving each variable either a fixed Boolean value or leaving it free. Let \(T|_\rho \) denote the tree obtained by collapsing every branch whose query variable is fixed by \(\rho \) into the corresponding subtree while keeping the unrestricted branches intact, and write \(\mathrm{depth}\) for the maximum length of a path from the root to a leaf. Then
Let \(T\) be a decision tree on \(n\) Boolean variables, let \(\rho \) be a restriction on those variables, and let \(x \in \{ \mathrm{true},\mathrm{false}\} ^n\) be an input. Then evaluating the restricted tree at \(x\) agrees with evaluating \(T\) at the assignment obtained by extending \(\rho \) with \(x\); that is, the restricted tree returns the same bit at \(x\) as \(T\) does at the point that takes each variable’s fixed value under \(\rho \) where \(\rho \) sets it and the value \(x_i\) on the free variables.
Let \(c\) be a Boolean circuit on \(m\) inputs whose depth is at least \(1\), and suppose each input coordinate \(i\) is fed by a Boolean function \(g_i \colon \{ 0,1\} ^n \to \{ 0,1\} \), so that \(c\) computes the map \(x \mapsto c\bigl(g_1(x),\dots ,g_m(x)\bigr)\). Fix a width bound \(l \ge 1\), and assume that every \(g_i\) is computed both by a DNF formula of width at most \(l\), each of whose terms contains at most one literal per variable, and by a CNF formula of width at most \(l\). Then there exist a number of inputs \(m'\), Boolean functions \(g'_1,\dots ,g'_{m'}\) each presented as a DNF formula of width at most \(l\) with the same restriction that no term contains two literals in the same variable, and a circuit \(c'\) of depth at most \(\mathrm{depth}(c) - 1\), such that \(c\bigl(g_1(x),\dots ,g_m(x)\bigr) = c'\bigl(g'_1(x),\dots ,g'_{m'}(x)\bigr)\) for every assignment \(x\).
Let \(c\) be a Boolean circuit on \(m\) inputs of depth exactly \(1\), and let \(g_1,\dots ,g_m \colon (\mathrm{Fin}\, n \to \mathbb {F}_2) \to \mathbb {F}_2\) be Boolean functions on \(n\) variables, each of which is computed by some DNF formula of width at most \(l\) and also by some CNF formula of width at most \(l\). Then the composed function \(x \mapsto c\big(g_1(x),\dots ,g_m(x)\big)\) — obtained by feeding \(g_i(x)\) into the \(i\)-th input of \(c\) — is, up to an overall sign, computed by a single DNF formula of width at most \(l\) each of whose terms has no repeated literal and mentions every variable at most once (in particular no term contains both a variable and its negation). Precisely, there exist such a DNF formula \(\varphi \) and a sign bit \(s\) so that, for every \(x\), \(\varphi (x) = c\big(g_1(x),\dots ,g_m(x)\big)\) when \(s\) is positive and \(\lnot \varphi (x) = c\big(g_1(x),\dots ,g_m(x)\big)\) when \(s\) is negative.
Let \(c\) be a Boolean circuit of depth exactly \(2\) on \(m\) inputs, and let \(g_1,\dots ,g_m \colon \{ 0,1\} ^n \to \{ 0,1\} \) be Boolean functions such that each \(g_i\) is computed both by a DNF formula of width at most \(l\) and by a CNF formula of width at most \(l\), where the width of a DNF or CNF formula is the maximum number of literals occurring in any single one of its terms or clauses. Then there exist a number \(m'\), DNF formulas \(D_1,\dots ,D_{m'}\) on \(n\) variables, and a Boolean circuit \(c'\) of depth at most \(1\) on \(m'\) inputs, such that each \(D_j\) has width at most \(l\), every term of every \(D_j\) contains each variable at most once (equivalently, its literals are pairwise distinct and no two of them share a variable), and for every assignment \(x \in \{ 0,1\} ^n\),
Let \(\psi \) be a CNF formula in the variables \(x_1,\dots ,x_n\), written as a conjunction of clauses, each clause a disjunction of literals. Then there is a CNF formula \(\psi '\) that computes the same Boolean function as \(\psi \), that is, \(\psi '(x)=\psi (x)\) for every assignment \(x\in \{ 0,1\} ^n\), whose width (the maximum number of literals in any of its clauses) is at most that of \(\psi \), and in which every clause is free of repetition in the strong sense that no literal occurs twice and no two distinct literals share a variable.
Fix a natural number \(m\) and layer sizes \(\ell _1, \dots , \ell _m \in \mathbb {N}\) whose total satisfies \(\sum _{i=1}^m \ell _i \le s\) for some \(s \in \mathbb {N}\). Let \(\alpha \ge 0\) and \(\beta \) be real numbers, and suppose \(a_1, \dots , a_m\) are real quantities with \(a_i \le \ell _i\, \alpha \) for each \(i\), together with a final real quantity \(b\) satisfying \(b \le \beta \). Then
Fix a natural number \(m\) and layer sizes \(\ell _0,\dots ,\ell _{m-1}\in \mathbb {N}\), and let \(s\in \mathbb {N}\) with \(s\gt 0\) bound their total, \(\sum _{i} \ell _i \le s\). Let \(\varepsilon \gt 0\) be a real number. Under these hypotheses the intended conclusion is that, on setting \(l=\log _2(2s/\varepsilon )\), the total weighted contribution is controlled by \(\varepsilon /2\), namely \(\sum _{i} \ell _i\, (1/2)^{l}\le \varepsilon /2\); however, the formal statement leaves both the value \(l\) and the conclusion as placeholders, so it asserts no mathematical claim beyond these hypotheses.
Let \(c\) be a Boolean circuit on \(n \ge 1\) variables of depth at most \(d\), size at most \(s\), and maximum fan-in at most \(w\), where \(d \ge 2\) and \(s, w, l \ge 1\), and fix a threshold \(t\). Apply a Bernoulli random restriction with parameter
so that each variable is left free with probability \(\delta \) and otherwise fixed to \(\mathtt{false}\) or \(\mathtt{true}\) with equal probability. Then the probability that the restricted function \(x \mapsto c(\rho .\mathrm{extend}\, x)\) has decision-tree depth exceeding \(t\) satisfies
as an upper bound.
Fix \(n \ge 1\) and let \(g_1, \dots , g_{s_2}\) be \(s_2 \ge 1\) DNF formulas on \(n\) variables (each a disjunction of terms, each term a conjunction of literals), each of width at most \(w \ge 1\), and such that within every term distinct literals involve distinct variables (no variable occurs more than once in a term). Let \(p\) satisfy \(0 \lt p \le 1\) and \(p \le \tfrac {1}{40w}\), and let \(\varepsilon \gt 0\) be such that \(\exp \! \bigl(-\tfrac {np}{3}\bigr) \le \tfrac {\varepsilon }{s_2}\). Then, under a Bernoulli\((p)\) random restriction \(\rho \), the probability that there exists an index \(i\) for which the restricted function \(g_i|_\rho \) admits no CNF formula of width at most \(l\) computing it (equal to it at every input) is at most
Let \(g_1,\dots ,g_{s_2}\) be a family of DNF formulas in \(n\) variables, and suppose \(n \ge 1\). Fix a width bound \(w \ge 1\) and assume each \(g_i\) has width at most \(w\); assume moreover that within every term of every \(g_i\) the literals are pairwise distinct and any two literals on the same variable coincide, so that no variable occurs more than once in a term. Let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le \frac{1}{40w}\) and \(p \le 1\), and draw a random restriction \(\rho \) on the \(n\) variables from the Bernoulli(\(p\)) distribution, in which each variable is independently left free with probability \(p\) and otherwise fixed to \(0\) or \(1\) each with probability \(\frac{1-p}{2}\). Then the probability that some gate \(g_i\) fails to be representable after restriction — that is, that there is an index \(i\) for which the restricted function \(g_i\! \restriction _\rho \) (the value \(x \mapsto g_i(\rho .\mathrm{extend}\, x)\)) is not computed by any CNF formula of width at most \(l\) — is at most
Let \(g_1,\dots ,g_m\) be DNF formulas on \(n\) Boolean variables, and let \(w\) be a positive integer such that every \(g_i\) has width at most \(w\); assume \(n \ge 1\), and that in each term the literals occupy distinct positions and any two literals sharing a variable are equal (so each variable is queried at most once per term). Let \(c\) be a Boolean circuit on the \(m\) gate-inputs whose depth is \(0\) (hence a single literal), and let \(F\) be the composed Boolean function on \(n\) variables given by \(F(x) = c\bigl(g_1(x),\dots ,g_m(x)\bigr)\). Fix natural numbers \(s\), \(l\), \(t\) with \(l \ge 1\), \(s \ge 1\), \(w \le l\), and \(\mathrm{size}(c) \le s\). Then, for a Bernoulli random restriction \(\rho \) with parameter \(\tfrac {1}{40l}\), the probability that the restriction of \(F\) along \(\rho \) — the function sending \(x\) to \(F\) evaluated with the coordinates fixed by \(\rho \) and the free coordinates read from \(x\) — has decision-tree depth exceeding \(t\) is at most
A bundle describing the second layer of a circuit over \(n\) variables: a number of gates, a family of DNF formulas indexed by the gates, a common width bound \(\mathrm{width}\gt 0\) such that every gate’s width is at most \(\mathrm{width}\), together with the hypothesis that within each term of each gate distinct literals have distinct variables.
Let \(g_1,\dots ,g_m \colon \{ 0,1\} ^n \to \{ 0,1\} \) be Boolean functions, and suppose that each \(g_i\) is computed both by a DNF formula of width at most \(l\) and by a CNF formula of width at most \(l\). Let \(C_1,\dots ,C_k\) be Boolean circuits on \(m\) inputs, each of depth at most \(1\). Then for every \(j \in \{ 1,\dots ,k\} \) there exist a DNF formula \(\varphi _j\) on the \(n\) variables and a sign \(\varepsilon _j \in \{ \mathrm{true},\mathrm{false}\} \) such that \(\varphi _j\) has width at most \(l\), each term of \(\varphi _j\) contains every variable at most once (so no term repeats a literal or carries a variable in both polarities), and, for every assignment \(x \in \{ 0,1\} ^n\), the value \(\varphi _j(x)\) when \(\varepsilon _j = \mathrm{true}\) and \(\lnot \varphi _j(x)\) when \(\varepsilon _j = \mathrm{false}\) equals \(C_j\bigl(g_1(x),\dots ,g_m(x)\bigr)\).
Flattens a list of lists \([\, \ell _1, \ell _2, \dots \, ]\) into the single list \(\ell _1 \mathbin {+\! \! +} \ell _2 \mathbin {+\! \! +} \cdots \), defined by recursion on the outer list. Used to merge the clause (or term) lists of several CNFs (or DNFs).
Given gate arrays \(g_1 : \mathrm{Fin}\, m_1 \to \alpha \) and \(g_2 : \mathrm{Fin}\, m_2 \to \alpha \), the merged array \(\mathrm{mergeGates}\, g_1\, g_2 : \mathrm{Fin}\, (m_1+m_2) \to \alpha \) takes its values from \(g_1\) on indices \(0,\dots ,m_1-1\) and from \(g_2\) on indices \(m_1,\dots ,m_1+m_2-1\).
Let \(g_1 : \mathrm{Fin}\, m_1 \to \alpha \) and \(g_2 : \mathrm{Fin}\, m_2 \to \alpha \) be two gate arrays, and let \(g_1 \ast g_2 : \mathrm{Fin}\, (m_1+m_2) \to \alpha \) denote their merge, which agrees with \(g_1\) on the indices \(0,\dots ,m_1-1\) and with \(g_2\) on the indices \(m_1,\dots ,m_1+m_2-1\). Then for every \(i \in \mathrm{Fin}\, m_1\), evaluating the merge at the index \(i\) regarded as an element of \(\mathrm{Fin}\, (m_1+m_2)\) (i.e. the same value \(i \lt m_1\)) returns \(g_1\, i\).
Fix \(n \in \mathbb {N}\), and let \(g_1\) and \(g_2\) be two arrays of DNF formulas over \(n\) variables, of lengths \(m_1\) and \(m_2\) respectively; write \(g_1 \sqcup g_2\) for their merge, the array of length \(m_1+m_2\) whose first \(m_1\) entries are the formulas of \(g_1\), in order, and whose remaining \(m_2\) entries are those of \(g_2\), in order. Suppose that every term occurring in any formula of \(g_1\), and every term occurring in any formula of \(g_2\), is a list of literals with no repetition. Then the same holds for the merged array: for every index \(k\) and every term \(t\) of the formula at position \(k\) in \(g_1 \sqcup g_2\), the literals of \(t\) are pairwise distinct.
Let \(g_1\) and \(g_2\) be two arrays of DNF formulas over \(n\) variables, indexed respectively by \(\{ 0,\dots ,m_1-1\} \) and \(\{ 0,\dots ,m_2-1\} \), and let \(g\) be their merge: the array indexed by \(\{ 0,\dots ,m_1+m_2-1\} \) that agrees with \(g_1\) on the first \(m_1\) indices and with \(g_2\), shifted down by \(m_1\), on the remaining \(m_2\). Suppose that in every formula of \(g_1\) and in every formula of \(g_2\), any two literals lying in one and the same term whose variables coincide are in fact the identical literal. Then the merged array \(g\) inherits this property: for every index \(k\) and any two literals occurring in a common term of the formula \(g(k)\) that share the same variable, the two literals are equal.
Let \(g_1\) and \(g_2\) be two families of DNF formulas over \(n\) variables, indexed respectively by \(m_1\) and \(m_2\) positions, and let \(l\) be a natural number. Suppose every formula in \(g_1\) and every formula in \(g_2\) has width at most \(l\). Then every formula in the merged family of \(m_1 + m_2\) formulas — those of \(g_1\) in order, followed by those of \(g_2\) — also has width at most \(l\).
Let \(g_1\) and \(g_2\) be families of DNF formulas over \(n\) variables, indexed by \(\mathrm{Fin}\, m_1\) and \(\mathrm{Fin}\, m_2\) respectively, and let \(\ell \) be a natural number. Suppose every formula \(g_1(k)\) has width at most \(\ell \). Form the merged array of length \(m_1+m_2\), which reproduces \(g_1\) on the first \(m_1\) indices and \(g_2\) on the remaining ones. Then for each \(i \in \mathrm{Fin}\, m_1\), the entry of the merged array at the position of \(i\) inside the first block again has width at most \(\ell \).
Let \(g_1\) and \(g_2\) be families of DNF formulas in \(n\) variables, indexed by \(\{ 0,\dots ,m_1-1\} \) and \(\{ 0,\dots ,m_2-1\} \) respectively, and let \(g\) be their merge: the family on \(\{ 0,\dots ,m_1+m_2-1\} \) that agrees with \(g_1\) on the first \(m_1\) indices and with \(g_2\) on the remaining ones, so that \(g(m_1+i)=g_2(i)\). Suppose \(l\) is a natural number for which every formula \(g_2(k)\) has width at most \(l\). Then for each index \(i\) of the second family, the merged formula \(g(m_1+i)\) has width at most \(l\).
Let \(\alpha \ge 0\), let \(s \in \mathbb {N}\), and consider \(m\) stages with layer sizes \(\ell _0, \ldots , \ell _{m-1} \in \mathbb {N}\) whose total does not exceed \(s\), i.e. \(\sum _{i} \ell _i \le s\). Suppose real numbers \(b_0, \ldots , b_{m-1}\) satisfy \(b_i \le \ell _i\, \alpha \) for every \(i\). Then \(\sum _{i} b_i \le s\, \alpha \).
Fix a Boolean flag \(\mathit{isAnd}\) and a finite list \(c_1,\dots ,c_k\) of Boolean circuits on \(m\) variables, and let \(d_1,\dots ,d_k\) be Boolean circuits on \(M\) variables, one for each position of the list. Given assignments \(g : \mathrm{Fin}\, m \to \mathbb {F}_2\) and \(g' : \mathrm{Fin}\, M \to \mathbb {F}_2\) such that every re-indexed child agrees with the corresponding original child, i.e. \(d_j\) evaluated at \(g'\) equals \(c_j\) evaluated at \(g\) for each position \(j\), the gate node with flag \(\mathit{isAnd}\) over \(c_1,\dots ,c_k\) evaluated at \(g\) equals the gate node with the same flag \(\mathit{isAnd}\) over \(d_1,\dots ,d_k\) evaluated at \(g'\). Here the node is the conjunction of its children’s values when \(\mathit{isAnd} = \mathrm{true}\) and their disjunction when \(\mathit{isAnd} = \mathrm{false}\).
Fix layer-2 data over \(n \ge 1\) variables: a family of \(m\) DNF gates \(G_1, \dots , G_m\) together with a common width bound \(w \gt 0\) satisfying \(\mathrm{width}(G_i) \le w\) for every \(i\). Let \(l \in \mathbb {N}\) and let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le \tfrac {1}{40\, w}\) and \(p \le 1\). Then, under a Bernoulli(\(p\)) random restriction \(\rho \) on the \(n\) variables, the probability that some gate fails the property — that is, the probability that it is *not* the case that for every \(i\) the restricted function \(x \mapsto G_i(\rho .\mathrm{extend}\, x)\) agrees, on all inputs, with some CNF formula of width at most \(l\) — is at most
Fix layer-2 data over \(n \ge 1\) Boolean variables: a family of \(N\) gates, each computing a DNF formula, with a common width bound \(w \ge 1\) so that every gate has width at most \(w\). Let \(l \in \mathbb {N}\) and let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le \tfrac {1}{40 w}\) and \(p \le 1\). Consider a Bernoulli(\(p\)) random restriction \(\rho \) on the \(n\) variables, in which each variable is independently left free with probability \(p\) and otherwise fixed to \(0\) or to \(1\), each with probability \(\tfrac {1-p}{2}\). Then the probability that some gate \(i\) has restricted function \(f_i\! \restriction _\rho \) (the gate’s DNF evaluated on the free coordinates, with the fixed coordinates set by \(\rho \)) of decision-tree depth greater than \(l\) satisfies
where \(\mathrm{depth}(g)\) denotes the least depth of a decision tree computing \(g\).
Let \(c\) be a Boolean circuit on \(n \ge 1\) variables with depth at most \(d\), size at most \(s\), and maximum fan-in at most \(w\), where \(d \ge 2\) and \(s, w, l \gt 0\). Fix a real number \(\varepsilon \gt 0\) and natural numbers \(l, t\) satisfying \(s \cdot \left(\tfrac {1}{2}\right)^{l} \le \varepsilon /2\) and \(\left(\tfrac {1}{2}\right)^{t} \le \varepsilon /2\), and put
Draw a random restriction \(\rho \) on the \(n\) variables by leaving each coordinate free with probability \(\delta \) and otherwise fixing it to a uniformly random Boolean value, all coordinates independent. Let \(c\! \restriction _{\rho }\) denote the restricted function of the free coordinates, obtained by evaluating \(c\) with the coordinates set by \(\rho \) held to their fixed values. Then the probability that \(c\! \restriction _{\rho }\) has decision-tree depth exceeding \(t\)—that is, that no decision tree of depth at most \(t\) computes it—satisfies
Fix an integer \(n \ge 1\), let \(w \ge 1\), and let \(g_1, \dots , g_{s_2}\) be \(s_2\) DNF formulas on \(n\) Boolean variables, each of width at most \(w\). Assume that in every term of every \(g_i\) no literal is repeated and no two distinct literals share a variable. Let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le \tfrac {1}{40w}\) (so in particular \(p \le 1\)). Then, under a Bernoulli(\(p\)) random restriction \(\rho \) on the \(n\) variables, the probability that every restricted formula \(x \mapsto g_i(\rho .\mathrm{extend}\, x)\) has decision-tree depth at most \(\ell \) is at least
Fix an integer \(w \ge 1\) and let \(D_1, \dots , D_{s_2}\) be DNF formulas over \(n \ge 1\) Boolean variables, each of width at most \(w\), and suppose that within every term of every \(D_i\) no variable occurs more than once (distinct literals of a term have distinct underlying variables, and no literal is repeated). Let \(0 \lt p \le \tfrac {1}{40w}\) (so in particular \(p \le 1\)), and draw a Bernoulli\((p)\) random restriction \(\rho \) on the \(n\) variables, in which each variable is independently left free with probability \(p\) and otherwise fixed to \(0\) or to \(1\), each with probability \(\tfrac {1-p}{2}\). Then for any target width \(l\), the probability that some restricted gate \(x \mapsto D_i(\rho .\mathrm{extend}\, x)\) fails to be computed by any CNF formula of width at most \(l\) is at most
Let \(g_1,\dots ,g_{s_2}\) be DNF formulas on \(n \ge 1\) variables, each of width at most \(w\), where \(w \ge 1\); assume that in every term of every \(g_i\) any two literals on the same variable coincide (so no term repeats a variable or contains both a literal and its negation) and no literal occurs twice in the term. Fix a natural number \(\ell \) and a real \(p\) with \(0 \lt p \le \tfrac {1}{40w}\) and \(p \le 1\). Then, for a Bernoulli-\(p\) random restriction \(\rho \) of the \(n\) variables,
where \(g_i|_\rho \) is the restriction of \(g_i\) by \(\rho \) and a CNF formula \(\psi \) *computes* \(g_i|_\rho \) when \(\psi \) and \(g_i|_\rho \) agree on every input.
Let \(c_1,\dots ,c_m\) be Boolean circuits on \(n\) variables, let \(\rho \) be a restriction of those variables, and let \(l\) be a natural number. For a Boolean function \(g\) write \(g\! \restriction _\rho \) for its restriction by \(\rho \), the function \(x \mapsto g(\rho .\mathrm{extend}\, x)\) that fixes each coordinate assigned by \(\rho \) and reads the remaining coordinates from \(x\). Suppose that for every child \(c\) among \(c_1,\dots ,c_m\) the restricted function \(c\! \restriction _\rho \) is computed by some decision tree of depth at most \(l\). Then there is a DNF formula \(\varphi \) of width at most \(l\) that computes the restriction \(F\! \restriction _\rho \) of the OR gate \(F\) over the children, i.e. \(\varphi (x) = F\! \restriction _\rho (x)\) for every assignment \(x\), where \(F\) evaluates to the disjunction of the values \(c_1,\dots ,c_m\).
Fix natural numbers \(m,n\) and a width bound \(l\). Let \(c_1,\dots ,c_k\) be Boolean circuits on \(m\) variables, each of which is a single literal, and let \(N\) be the OR node whose children are exactly these circuits. Suppose that to each variable index \(i \in \mathrm{Fin}\, m\) there is associated a gate function \(g_i\) taking a Boolean assignment \(x : \mathrm{Fin}\, n \to \{ 0,1\} \) to a Boolean value, and that every \(g_i\) is computed both by some DNF formula on the \(n\) inputs of width at most \(l\) and by some CNF formula on the \(n\) inputs of width at most \(l\). Then the composite function \(x \mapsto N\bigl((g_i(x))_{i}\bigr)\) — the disjunction of the literal children evaluated at the gate values \(g_i(x)\) — is itself computed by a DNF formula \(\varphi \) on the \(n\) inputs of width at most \(l\), in the sense that \(\varphi (x)\) equals the value of \(N\) on the assignment \(i \mapsto g_i(x)\) for every \(x\).
Fix a number of variables \(n\) and let assignments \(x\) range over \(\{ 0,1\} ^n\). Let \(C_1,\dots ,C_r\) be Boolean circuits, each on the same set of \(m\) inputs, and let \(\gamma _1,\dots ,\gamma _m\) be functions assigning to every assignment \(x\) a bit, so that feeding input \(i\) the value \(\gamma _i(x)\) turns each \(C_j\) into a Boolean function of \(x\). Fix a width bound \(l\) and, for each \(j\), a depth bound \(b_j\). Assume every child admits a DNF reduction: for each \(j\) there are a Boolean circuit \(c_j\) of depth at most \(b_j\) on some number \(m_j\) of inputs, together with DNF formulas \(g_{j,1},\dots ,g_{j,m_j}\) over \(x_1,\dots ,x_n\), each of width at most \(l\) and each of whose terms has its literals on pairwise distinct variables, such that
Then these per-child families can be amalgamated into a single one: there exist a family of DNF formulas \(h_1,\dots ,h_M\) over \(x_1,\dots ,x_n\), each of width at most \(l\) and each with all its terms’ literals on pairwise distinct variables, together with circuits \(C'_1,\dots ,C'_r\) on these \(M\) shared inputs, such that \(\mathrm{depth}(C'_j) \le b_j\) and
Let \(c\) be a Boolean circuit on \(m_1\) variables, and let \(g_1 : \mathrm{Fin}\, m_1 \to \mathbb {F}_2\) and \(g_2 : \mathrm{Fin}\, m_2 \to \mathbb {F}_2\) be assignments. Write \(\iota : \mathrm{Fin}\, m_1 \to \mathrm{Fin}\, (m_1+m_2)\), \(i \mapsto i\), for the order-preserving inclusion of the first block, and let \(g\) be the merged assignment on \(m_1 + m_2\) variables that agrees with \(g_1\) on the first \(m_1\) indices and with \(g_2\) on the last \(m_2\). Then the circuit obtained by replacing each literal index \(i\) of \(c\) by \(\iota (i)\) (keeping its sign), evaluated under \(g\), has the same value as \(c\) evaluated under \(g_1\).
Let \(m_1, m_2\) be natural numbers, let \(c\) be a Boolean circuit on \(m_2\) variables, and let \(g_1 : \mathbb {F}_2^{m_1}\) and \(g_2 : \mathbb {F}_2^{m_2}\) be Boolean assignments. Form the merged assignment on \(m_1 + m_2\) variables that takes the values of \(g_1\) on indices \(0,\dots ,m_1-1\) and the values of \(g_2\) on indices \(m_1,\dots ,m_1+m_2-1\), and re-index \(c\) by the right embedding \(i \mapsto m_1 + i\) from \(\mathrm{Fin}\, m_2\) into \(\mathrm{Fin}\, (m_1+m_2)\). Then
where each side denotes circuit evaluation; that is, evaluating the re-indexed circuit under the merged assignment gives the same Boolean value as evaluating \(c\) directly under \(g_2\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function, let \(\rho \) be a restriction on the \(n\) variables (fixing some coordinates to Boolean values and leaving the rest free), and write \(f\! \restriction _\rho \) for the restricted function \(x \mapsto f(\rho .\mathrm{extend}\, x)\), obtained by evaluating \(f\) with the fixed coordinates set by \(\rho \) and the free coordinates supplied by \(x\). If \(f\! \restriction _\rho \) has decision-tree depth at most \(l\), then there is a CNF formula \(\psi \) whose every clause contains at most \(l\) literals and which computes \(f\! \restriction _\rho \) on every input, that is, \(\psi (x) = f\! \restriction _\rho (x)\) for all \(x \in \{ 0,1\} ^n\).
Let \(f : (\mathbb {F}_2^{\, n} \to \{ 0,1\} )\) be a Boolean function on \(n\) variables and let \(\rho \) be a restriction that fixes some variables to Boolean values and leaves the others free, so that the restricted function \(f|_\rho \) sends a point \(x\) to the value of \(f\) on the assignment agreeing with \(\rho \) on the fixed coordinates and with \(x\) on the free ones. If \(f|_\rho \) has decision-tree depth at most \(l\) — that is, some decision tree of depth at most \(l\) computes it — then there exists a DNF formula \(\varphi \) whose width (the maximum number of literals in any of its terms) is at most \(l\) and which agrees with \(f|_\rho \) on every input.
Let \(f : (\mathrm{Fin}\, n \to \mathbb {F}_2) \to \mathbb {F}_2\) be a Boolean function on \(n\) variables, and let \(\rho _1, \rho _2 : \mathrm{Fin}\, n \to \{ 0, 1, *\} \) be two restrictions, where a coordinate is either fixed to a Boolean value or left free. Write \(\rho _1 * \rho _2\) for their composition, the restriction that on each coordinate takes the value fixed by \(\rho _1\) if there is one and otherwise the value assigned by \(\rho _2\), and for a restriction \(\rho \) write \(f\! \restriction _\rho \) for the function \(x \mapsto f(\rho .\mathrm{extend}\, x)\) obtained by evaluating \(f\) with the fixed coordinates of \(\rho \) substituted and the free coordinates read off from \(x\). Then restricting \(f\) by the composite restriction agrees with restricting first by \(\rho _1\) and then by \(\rho _2\):
Let \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be a Boolean function on \(n\) variables, and let \(\rho _1, \rho _2\) be two restrictions on those variables, each fixing some coordinates to Boolean values and leaving the rest free. Write \(\rho _1 \circ \rho _2\) for their composition, the restriction that on each coordinate uses the value fixed by \(\rho _1\) when there is one and otherwise the value assigned by \(\rho _2\), and for a restriction \(\rho \) write \(f\! \restriction _\rho \) for the function obtained from \(f\) by fixing the coordinates set by \(\rho \) and evaluating on the remaining free coordinates. Then restricting \(f\) by the composition \(\rho _1 \circ \rho _2\) gives the same function as first restricting \(f\) by \(\rho _1\) and then restricting the function \(f\! \restriction _{\rho _1}\) by \(\rho _2\):
Let \(f, g : (\mathrm{Fin}\, n \to \mathbb {F}_2\text{-valued Boolean inputs}) \to \mathrm{Bool}\) be two Boolean functions on \(n\) variables that agree everywhere, \(f(y) = g(y)\) for every input \(y\), and let \(\rho \) be a restriction on \(n\) variables. Then the restricted functions agree as well: writing \(f_\rho (x) = f(\rho .\mathrm{extend}\, x)\) for the restriction of \(f\) by \(\rho \), one has \(f_\rho (x) = g_\rho (x)\) for every \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\).
Let \(f, g : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be two Boolean functions that agree at every point, so that \(f(x) = g(x)\) for all \(x\). Then for every restriction \(\rho \) on \(n\) variables, the restricted functions coincide: the restriction of \(f\) by \(\rho \) equals the restriction of \(g\) by \(\rho \).
Let \(f_1, \dots , f_k\) be Boolean functions on \(n\) variables and let \(\rho \) be a restriction of those variables, and write \(f|_\rho \) for the restriction of a function \(f\) along \(\rho \) (the function \(x \mapsto f(\rho .\mathrm{extend}\, x)\)). Then restricting the pointwise conjunction \(f_1 \wedge \cdots \wedge f_k\) along \(\rho \) agrees, at every point \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\), with the pointwise conjunction of the individually restricted functions \(f_1|_\rho , \dots , f_k|_\rho \).
Fix \(n\) variables and let \(C_1,\dots ,C_k\) be a finite list of Boolean circuits on them, let \(\rho \) be a restriction, and let \(b\) be a Boolean flag. For a Boolean function \(f\) write \(f_\rho \) for its restriction by \(\rho \), the function \(x \mapsto f(\rho .\mathrm{extend}\, x)\). Form the circuit node with children \(C_1,\dots ,C_k\), read as an AND gate when \(b\) is true and as an OR gate when \(b\) is false, and let \(g\) be its evaluation. Then the restriction of the gate’s evaluation agrees, as a Boolean function, with the gate applied to the restricted evaluations of the children: for every assignment \(x\),
where \((C_i)_\rho \) denotes the restriction by \(\rho \) of the evaluation of \(C_i\), an empty conjunction is \(\mathrm{true}\), and an empty disjunction is \(\mathrm{false}\).
Fix real parameters \(p, q\) with \(0 \lt p \le 1\) and \(0 \lt q \le 1\), and let \(E\) be any decidable predicate on the restrictions of \(n\) variables. Write \(w_p(\rho )\) for the Bernoulli weight of a restriction \(\rho \), write \(\Pr _p[E]\) for the Bernoulli(\(p\)) probability of \(E\), and let \(\rho _1 * \rho _2\) denote the composition of \(\rho _1\) and \(\rho _2\). Then
where the sum runs over all restrictions \(\rho _1\) on the \(n\) variables, and the inner probability is that of the event that \(E\) holds of the composition \(\rho _1 * \rho _2\) as \(\rho _2\) ranges over a Bernoulli(\(q\)) restriction.
Fix \(n\), let \(E\) be an event on restrictions of \(n\) variables, and let \(0 \lt p \le 1\) and \(0 \lt q \le 1\) be real parameters. Then the probability that \(E\) holds under a Bernoulli\((pq)\) random restriction is at most the probability, under a Bernoulli\((p)\) random restriction \(\rho _1\), that there is some restriction \(\rho _2\) for which \(E\) holds of the composition of \(\rho _1\) with \(\rho _2\):
Let \(s\) be a positive integer and let \(\varepsilon \gt 0\) be a real number. Then
where the exponent \(\log _2(2s/\varepsilon )\) is the base-\(2\) logarithm.
Fix an integer \(n \ge 1\) and a positive integer \(l\), and let \(\phi _1, \dots , \phi _s\) be DNF formulas on \(n\) Boolean variables, each of width at most \(l\) and each written so that no variable occurs more than once in any single term — that is, no term repeats a literal, and no term contains two literals on the same variable. Let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le \tfrac {1}{40 l}\) (so in particular \(p \le 1\)). Then a Bernoulli\((p)\) random restriction \(\rho \) on the \(n\) variables satisfies, with probability at least
that for every \(i\) the restricted function \(x \mapsto \phi _i\) evaluated on the coordinates left free by \(\rho \) (with the fixed coordinates set by \(\rho \)) has decision-tree depth at most \(l\).
Fix positive integers \(n\) and \(l\), and let \(g_1,\dots ,g_{s_i}\) be DNF formulas on \(n\) variables, each of width at most \(l\), and each in reduced form: no term repeats a literal, and within a term any two literals on the same variable coincide, so every variable occurs at most once in each term. Let \(p\) be a real parameter with \(0 \lt p \le 1\) and \(p \le \tfrac {1}{40\, l}\). Draw a Bernoulli(\(p\)) random restriction \(\rho \) of the \(n\) variables, which independently leaves each variable free with probability \(p\) and otherwise fixes it to \(0\) or \(1\) each with probability \((1-p)/2\). Then the probability that some restricted gate fails to be computed by a narrow CNF — that is, for some \(i\) there is no CNF formula \(\psi \) of width at most \(l\) satisfying \(\psi (x) = g_i(\rho .\mathrm{extend}\, x)\) for every assignment \(x\) — is at most
Let \(n\), \(m\), and \(l\) be natural numbers, let \(\rho _1\) be a restriction on \(n\) variables, and let \(g_1, \dots , g_m\) be DNF formulas on \(n\) variables. For each \(i\), write \(g_i|_{\rho _1}\) for the restricted Boolean function \(x \mapsto g_i\bigl(\rho _1.\mathrm{extend}\, x\bigr)\), that is, \(g_i\) evaluated on the input obtained by setting the fixed coordinates of \(\rho _1\) and reading the free coordinates from \(x\). Suppose that for every \(i\) the decision-tree depth of \(g_i|_{\rho _1}\) is at most \(l\). Then there exist DNF formulas \(g_1', \dots , g_m'\) such that, for every \(i\): the width of \(g_i'\) is at most \(l\); \(g_i'\) computes \(g_i|_{\rho _1}\) on every input; within each term of \(g_i'\) any two literals on the same variable are identical; and no term of \(g_i'\) contains a repeated literal.
Let \(g_1,\dots ,g_m\) be DNF formulas on \(n\) Boolean variables, let \(\rho \) be a restriction of these variables, and let \(l\in \mathbb {N}\). For each \(i\), write \(f_i\) for the restriction of \(g_i\) by \(\rho \), that is, the Boolean function \(x \mapsto g_i(\rho .\mathrm{extend}\, x)\) that evaluates \(g_i\) with the coordinates fixed by \(\rho \) set accordingly and the free coordinates supplied by \(x\), and suppose that every \(f_i\) has decision-tree depth at most \(l\). Then for each \(i\) there is a DNF formula \(\varphi _i\) of width at most \(l\) that computes \(f_i\), in which no term repeats a literal and no two literals of the same term share a variable; and there is a CNF formula \(\psi _i\) of width at most \(l\) that also computes \(f_i\).
Let \(f\) be a CNF formula in \(n \ge 1\) Boolean variables whose width is at most \(w\), with \(w \ge 1\), and let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le 1/(40w)\) and \(p \le 1\). Consider a Bernoulli(\(p\)) random restriction \(\rho \) of the \(n\) variables, in which each variable is independently left free with probability \(p\) and otherwise fixed to \(\mathtt{false}\) or to \(\mathtt{true}\), each with probability \((1-p)/2\), and write \(f_\rho \) for the Boolean function on the free coordinates obtained by evaluating \(f\) with the fixed coordinates set by \(\rho \). Then for every \(t \in \mathbb {N}\),
where \(\mathrm{depth}(f_\rho )\) denotes the decision-tree depth of \(f_\rho \).
Let \(f\) be a DNF formula on \(n \ge 1\) Boolean variables whose width is at most \(w\), where \(w \ge 1\), and let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le 1\) together with \(p \le \tfrac {1}{40w}\). Draw a Bernoulli(\(p\)) random restriction \(\rho \), which independently leaves each of the \(n\) variables free with probability \(p\) and otherwise fixes it to a uniformly random Boolean value, and let \(f_\rho \) denote the function of the free variables obtained by evaluating \(f\) with the fixed coordinates set according to \(\rho \). Then for every \(t \in \mathbb {N}\),
where \(\operatorname {depth}(f_\rho )\) is the decision-tree depth of \(f_\rho \), i.e. the least depth of a decision tree computing it.
Fix an integer \(n \ge 1\) and a Boolean function \(f : \{ 0,1\} ^n \to \{ 0,1\} \) whose decision-tree depth is at most \(w\), for some integer \(w \ge 1\). Let \(p\) be a real parameter with \(0 \lt p \le 1/(40w)\) (so in particular \(p \le 1\)), and draw a Bernoulli\((p)\) random restriction \(\rho \) on the \(n\) variables: independently, each variable is left free with probability \(p\) and otherwise fixed to \(0\) or to \(1\), each with probability \((1-p)/2\). Then for every \(t \in \mathbb {N}\), the restricted function \(f\restriction \rho \) (that is, \(f\) evaluated on the free coordinates with the remaining coordinates set by \(\rho \)) satisfies
Let \(g\) be a DNF formula on \(n \ge 1\) variables whose width is at most \(w\), where \(w \ge 1\), and suppose that within each term of \(g\) any two literals having the same variable are equal and no literal is repeated. Let \(l\) be a natural number, and let \(p\) be a real number with \(0 \lt p \le 1\) and \(p \le \tfrac {1}{40w}\). Then, under a Bernoulli\((p)\) random restriction \(\rho \), the probability that the restricted function \(g|_\rho \) is computed by no CNF formula of width at most \(l\) is at most
Let \(g\) be a CNF formula on \(n \ge 1\) variables whose width is at most \(w\), where \(w \ge 1\), and suppose that within each clause no two distinct literals share a variable and no literal is repeated. Fix a natural number \(l\) and a real parameter \(p\) with \(0 \lt p \le \tfrac {1}{40w}\), so in particular \(p \le 1\). Then, drawing a Bernoulli\((p)\) random restriction \(\rho \) on the \(n\) variables, the probability that the restricted function \(g|_\rho \) (the function \(x \mapsto g(\rho .\mathrm{extend}\, x)\) on the free coordinates) cannot be computed exactly by any DNF formula of width at most \(l\) is at most
Let \(D_0,\dots ,D_{s-1}\) be \(s\) DNF formulas on \(n\) Boolean variables, with \(n \ge 1\), each of width at most \(w\) where \(w \ge 1\), and suppose that in every term of every \(D_i\) each variable occurs at most once (no term repeats a literal, and any two literals of a term on the same variable coincide). Let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le \tfrac {1}{40w}\) and \(p \le 1\). Draw a Bernoulli\((p)\) random restriction \(\rho \), which independently leaves each variable free with probability \(p\) and otherwise fixes it to \(\mathtt{false}\) or \(\mathtt{true}\) each with probability \((1-p)/2\). Then the probability that some restricted formula \(D_i\! \restriction _\rho \) computes a Boolean function of decision-tree depth greater than \(l\) is at most
Fix a number of variables \(m\), a gate type \(isAnd \in \{ \mathrm{true}, \mathrm{false}\} \), and a list \(cs\) of Boolean circuits on \(m\) variables, and form the circuit \(\mathrm{node}(isAnd, cs)\) whose root gate is applied to the children in \(cs\). If this circuit has depth at most \(1\), then every child \(c\) in \(cs\) is a literal; that is, there exists a literal \(l\) with \(c = \mathrm{lit}(l)\).
Let \(C\) be the Boolean circuit consisting of a single gate (an AND gate or an OR gate, according to a flag \(\mathit{isAnd}\)) applied to a list \(cs\) of child circuits, so that the depth of \(C\) equals \(1\) plus the maximum depth among the circuits in \(cs\). If the depth of \(C\) is at most \(1\), then every child circuit \(c \in cs\) has depth \(0\); that is, each child is a literal.
Let \(c\) be a Boolean circuit whose depth is at least \(1\). Then \(c\) is not a literal: it is an internal gate, so there exist a Boolean flag \(\mathit{isAnd}\) and a list \(cs\) of child circuits for which \(c\) is the node with gate type \(\mathit{isAnd}\) applied to the children \(cs\).
Let \(p_1, p_2 \in (0,1]\), let \(E\) and \(A\) be events on restrictions of \(n\) variables, and let \(\beta \ge 0\). Write \(\Pr _p[\, \cdot \, ]\) for the probability of an event under a Bernoulli(\(p\)) random restriction, and for restrictions \(\rho _1, \rho _2\) let \(\rho _1 \circ \rho _2\) denote their composition (the values fixed by \(\rho _1\), filled in by \(\rho _2\) on the coordinates \(\rho _1\) leaves free). Suppose that for every restriction \(\rho _1\) for which \(A(\rho _1)\) fails, one has \(\Pr _{p_2}[\, E(\rho _1 \circ \rho _2)\, ] \le \beta \). Then
Let \(p_1, p_2 \in (0,1]\), let \(E\) and \(A\) be predicates on restrictions of \(n\) variables, and let \(\beta \ge 0\). For a parameter \(p\) write \(\mathbb {P}_p[\, \cdot \, ]\) for the Bernoulli event probability under a Bernoulli(\(p\)) random restriction, and for two restrictions \(\rho _1, \rho _2\) write \(\rho _1 \ast \rho _2\) for their composition (using the value \(\rho _1\) fixes on each coordinate, and otherwise the value from \(\rho _2\)). Suppose that for every restriction \(\rho _1\) with \(\neg A(\rho _1)\) one has \(\mathbb {P}_{p_2}\! \left[\, E(\rho _1 \ast \rho _2) \, \right] \le \beta \). Then
Fix \(n\), and let \(p,q\) be real parameters with \(0\le p\le 1\) and \(0\le q\le 1\). For a parameter \(r\) write \(w_r(\rho )\) for the Bernoulli weight of a restriction \(\rho \) on \(n\) variables at parameter \(r\), and write \(\mathbb {P}_r[E]\) for the Bernoulli event probability of a predicate \(E\) on restrictions at parameter \(r\). Then for any two predicates \(A\) and \(B\) on restrictions of \(n\) variables,
where the inner sum ranges over all restrictions \(\rho _1\) on \(n\) variables and, for each fixed \(\rho _1\), the predicate \(\rho _2 \mapsto A(\rho _1)\lor B(\rho _2)\) holds at a second-stage restriction \(\rho _2\) exactly when \(A\) holds at \(\rho _1\) or \(B\) holds at \(\rho _2\).
Regard each coordinate of a random restriction as being in one of three states: *free* (written \(\ast \)) or *fixed* to one of the two Boolean values. For a parameter \(p \in \mathbb {R}\), let \(w_p\) be the per-variable weight assigning \(w_p(\ast ) = p\) to the free state and \((1-p)/2\) to each of the two fixed states, and define the composite \(a \circ b\) of two states to equal \(a\) when \(a\) is fixed and \(b\) when \(a\) is free. Then for all \(p, q \in \mathbb {R}\) and every state \(c\),
where both sums range over the three states and \([\, \cdot \, ]\) equals \(1\) when the enclosed equation holds and \(0\) otherwise.
Let \(t_1, \dots , t_k\) be a finite list of terms, each a conjunction of literals, and let \(\ell \) be a natural number; write \(w(t)\) for the width of a term (its number of literals). Then the maximum width over the list, taken to be \(0\) when the list is empty, satisfies \(\max _{1 \le i \le k} w(t_i) \le \ell \) if and only if \(w(t) \le \ell \) for every term \(t\) in the list.
Fix \(N \ge 1\) experts and \(T\) rounds, a learning rate \(\eta \gt 0\), and a valid loss sequence \(\ell \), so that \(0 \le \ell _t(i) \le 1\) for every round \(t\) and expert \(i\). Writing \(W_t = \sum _{i=1}^N \exp \! \bigl(-\eta L_t(i)\bigr)\) for the Hedge potential and \(\widehat{\ell }_t = \sum _{i=1}^N p_t(i)\, \ell _t(i)\) for the expected loss under the Hedge distribution, the one-round change in log-potential is controlled by the expected loss: for each round \(t\),
Fix \(N \ge 1\) experts and a learning rate \(\eta \gt 0\), and let \(\ell \) be a valid loss sequence over \(T\) rounds, so that every individual loss satisfies \(0 \le \ell _t(i) \le 1\). Write \(W_t = \sum _{i=1}^{N} \exp (-\eta L_t(i))\) for the Hedge potential after \(t\) rounds, and \(\widehat{\ell }_t\) for the expected loss of the learner at round \(t\) under the Hedge distribution. Then for every round \(t\) (with \(0 \le t \lt T\)),
Let \(\eta \in \mathbb {R}\) and let \(\ell \) be a loss sequence for \(N\) experts over \(T\) rounds, with potential \(W_t = \texttt{potential}\ \eta \ \ell \ t\). Then the one-step changes of the log potential telescope:
\(\texttt{LossHistory}\ N\ t\) is the type of length-\(t\) histories of loss vectors for \(N\) experts: it is the function type \(\mathrm{Fin}\, t \to \mathrm{Fin}\, N \to \mathbb {R}\), recording all loss vectors that have been revealed in the past \(t\) rounds.
Given a history \(h : \texttt{LossHistory}\ N\ t\), \(\text{\texttt{LossHistory.toLossSeq}}\ h\) coerces it to a \(\texttt{LossSeq}\ N\ t\), enabling the reuse of definitions and theorems from the pathwise Hedge formalization that are stated over finite loss sequences.
A loss sequence \(\ell \) for \(N\) experts over \(T\) rounds is a function \(\ell : \mathrm{Fin}\, T \to \mathrm{Fin}\, N \to \mathbb {R}\), i.e., at each round \(t \lt T\) the adversary reveals a loss vector \(\ell _t : \mathrm{Fin}\, N \to \mathbb {R}\) whose values are intended to lie in \([0,1]\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function, let \(d \ge 1\), and let \(\varepsilon \) be a real number with \(0 \le \varepsilon \le 1\). Suppose \(f\) is \(\varepsilon \)-far from degree \(\le d\), meaning that \(\operatorname{dist}(f,g) \ge \varepsilon \) for every Boolean function \(g\) of degree \(\le d\). Then, writing \(\widehat{(-1)^f}(S)\) for the Fourier coefficient at \(S\) of the \(\pm 1\) lift \(x \mapsto (-1)^{f(x)}\), we have
Fix natural numbers \(n\) and \(k\), and let \(\{ 0,1\} ^n\) denote the Boolean hypercube of dimension \(n\). The set of \(k\)-tuples of points of \(\{ 0,1\} ^n\), that is \(\left(\{ 0,1\} ^n\right)^k\), has exactly \(2^{nk}\) elements.
For every subset \(S \subseteq [n]\) and every integer \(d \ge 1\), the Walsh–Fourier character \(\chi _S\) has degree \(\le d\): for every base point \(x \in \{ 0,1\} ^n\) and all direction vectors \(h_1, \ldots , h_{d+1} \in \{ 0,1\} ^n\), the \((d{+}1)\)-fold Gowers product satisfies \(\mathrm{GP}(\chi _S,\, d{+}1,\, x,\, h_1, \ldots , h_{d+1}) = 1\).
For every subset \(S \subseteq [n]\), the Walsh–Fourier character \(\chi _S\) has degree at most \(1\): every second-order Gowers product of \(\chi _S\) equals \(1\). Explicitly, for all base points \(x \in \{ 0,1\} ^n\) and all direction vectors \(h_1, h_2 \in \{ 0,1\} ^n\),
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function of degree at most \(1\); that is, its \(\pm 1\) lift \(x \mapsto (-1)^{f(x)}\) has degree at most \(1\), meaning every second-order multiplicative derivative is trivial: for all \(x, h_1, h_2 \in \{ 0,1\} ^n\) the order-\(2\) Gowers product of \((-1)^f\) at \(x\) in directions \(h_1, h_2\) equals \(1\). Then \(f\) is affine over \(\mathbb {F}_2\): either \(f\) itself is linear, satisfying \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y\), or its pointwise negation \(x \mapsto \lnot f(x)\) is linear.
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\), and recall that \(f\) has degree \(\le d\) when all of its \((d{+}1)\)-fold multiplicative derivatives vanish, i.e. the Gowers product \(\mathrm{GP}(f, d{+}1, x, h_1,\ldots ,h_{d+1}) = 1\) for every base point \(x \in \{ 0,1\} ^n\) and all directions \(h_1,\ldots ,h_{d+1} \in \{ 0,1\} ^n\). If \(f\) has degree \(\le d\), then \(f\) also has degree \(\le d+1\); that is, \(\mathrm{GP}(f, d{+}2, x, h_1,\ldots ,h_{d+2}) = 1\) for every \(x\) and all directions \(h_1,\ldots ,h_{d+2} \in \{ 0,1\} ^n\).
The probability that the \((d{+}1)\)-fold derivative test accepts \(f\):
Fix \(n\) and an integer \(d \ge 0\), and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function with \(\pm 1\) lift \((-1)^f : \{ 0,1\} ^n \to \{ \pm 1\} \). Then the probability that the \((d+1)\)-fold derivative low-degree test accepts \(f\) is
where \(\| (-1)^f\| _{U^{d+1}}^{2^{d+1}}\) is the \((d+1)\)-st Gowers uniformity norm of the lift raised to the power \(2^{d+1}\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function of degree at most \(d\), meaning that its \(\pm 1\) lift \((-1)^f\) has every \((d{+}1)\)-fold Gowers product equal to \(1\): for all base points \(x\) and all direction vectors \(h_1,\ldots ,h_{d+1} \in \{ 0,1\} ^n\), one has \(\mathrm{GP}\bigl((-1)^f, d{+}1, x, h_1,\ldots ,h_{d+1}\bigr) = 1\). Then the \((d{+}1)\)-fold derivative test accepts \(f\) with probability exactly \(1\); that is,
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function that does not have degree at most \(d\), meaning that its \(\pm 1\) lift \((-1)^f\) has at least one nonvanishing \((d{+}1)\)-fold multiplicative derivative: there exist a base point \(x\) and directions \(h_1,\ldots ,h_{d+1} \in \{ 0,1\} ^n\) with \(\mathrm{GP}\! \left((-1)^f,\, d{+}1,\, x,\, h_1,\ldots ,h_{d+1}\right) \neq 1\). Then the \((d{+}1)\)-fold derivative test accepts \(f\) with probability strictly less than \(1\).
Fix \(d \ge 1\) and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function that is \(\varepsilon \)-far from degree \(\le d\); that is, \(0 \le \varepsilon \le 1\) and every Boolean function \(g\) of degree \(\le d\) satisfies \(\operatorname{dist}(f,g) \ge \varepsilon \). Then the \((d{+}1)\)-fold derivative test accepts \(f\) with probability at most \(1 - \varepsilon \):
where \(x\) and the direction vectors \(h_1,\ldots ,h_{d+1}\) range uniformly over \(\{ 0,1\} ^n\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function that is \(\varepsilon \)-far from degree \(\le d\), where \(d \ge 2\); in particular \(0 \le \varepsilon \le 1\). For a shift \(h \in \{ 0,1\} ^n\), write \(\partial _h f\) for the discrete derivative \(\partial _h f(x) = f(x) \oplus f(x \oplus h)\). Then there exists a real number \(\delta \ge \varepsilon \) such that for every \(h \in \{ 0,1\} ^n\) there is a Boolean function \(g : \{ 0,1\} ^n \to \{ 0,1\} \) for which, if \(g\) has degree \(\le d-1\), then \(\operatorname{dist}(\partial _h f, g) \ge \delta \).
Fix \(n\) and an integer \(d \ge 1\), and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function. If \(f\) is \(\varepsilon \)-far from degree \(\le d\) — that is, \(0 \le \varepsilon \le 1\) and every Boolean function \(g\) of degree \(\le d\) satisfies \(\operatorname{dist}(f,g) \ge \varepsilon \) — then \(\varepsilon \le 1/2\).
Fix a Boolean function \(f : \{ 0,1\} ^n \to \{ 0,1\} \), a real number \(\varepsilon \) with \(0 \le \varepsilon \le 1\), and natural numbers \(d' \le d\). If \(f\) is \(\varepsilon \)-far from degree \(\le d\) — that is, \(\operatorname{dist}(f, g) \ge \varepsilon \) for every Boolean function \(g\) of degree \(\le d\) — then \(f\) is also \(\varepsilon \)-far from degree \(\le d'\).
Let \(f:\{ 0,1\} ^n\to \{ 0,1\} \) be a Boolean function, let \(d\ge 1\), and let \(\varepsilon \) satisfy \(0\le \varepsilon \le 1\). Suppose \(f\) is \(\varepsilon \)-far from degree \(\le d\); that is, \(\Pr _{x}[f(x)\ne g(x)]\ge \varepsilon \) for every Boolean function \(g:\{ 0,1\} ^n\to \{ 0,1\} \) of degree \(\le d\). Writing \((-1)^f:\{ 0,1\} ^n\to \{ \pm 1\} \) for the \(\pm 1\) lift of \(f\) and \(\widehat{(-1)^f}(S)=\mathbb {E}_x[(-1)^{f(x)}\chi _S(x)]\) for its Fourier coefficient at \(S\), we then have, for every \(S\subseteq [n]\),
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function and let \(d\) be a nonnegative integer, and write \((-1)^f : \{ 0,1\} ^n \to \{ \pm 1\} \) for its sign lift. Then the \(2^{d+1}\)-th power of the order-\((d+1)\) Gowers uniformity norm of \((-1)^f\) equals \(1\),
if and only if \((-1)^f\) has degree at most \(d\), meaning that for every base point \(x \in \{ 0,1\} ^n\) and all direction vectors \(h_1,\ldots ,h_{d+1} \in \{ 0,1\} ^n\) the \((d+1)\)-fold Gowers product \(\mathrm{GP}\big((-1)^f, d{+}1, x, h_1,\ldots ,h_{d+1}\big)\) equals \(1\).
Let \(d \ge 1\) and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function that is \(\varepsilon \)-far from degree \(\le d\), meaning \(0 \le \varepsilon \le 1\) and \(\operatorname{dist}(f, g) \ge \varepsilon \) for every Boolean function \(g\) of degree \(\le d\). Then there is a constant \(c \gt 0\) for which the \(\pm 1\) lift \((-1)^{f}\) satisfies
Let \(d \ge 1\) and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function, and write \(F = (-1)^f\) for its \(\pm 1\) lift. Suppose \(f\) is \(\varepsilon \)-far from degree \(\le d\): that is, \(0 \le \varepsilon \le 1\), and every Boolean function \(g\) of degree \(\le d\) satisfies \(\operatorname{dist}(f,g) \ge \varepsilon \), where \(\operatorname{dist}(f,g) = \Pr _{x}[f(x) \ne g(x)]\) over a uniform \(x \in \{ 0,1\} ^n\). Then the second Gowers uniformity norm of \(F\) obeys
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function on the hypercube, and let \(g = (-1)^f\) be its \(\pm 1\) lift, the real-valued function \(g(x) = (-1)^{f(x)} \in \{ \pm 1\} \). Then for every \(k \in \mathbb {N}\), the \(k\)-th Gowers uniformity norm of \(g\) (in power form) satisfies
Fix \(n, d \in \mathbb {N}\) and let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions. If \(g\) has degree \(\le d\)—that is, every \((d{+}1)\)-fold multiplicative derivative of \(g\) equals \(1\)—then multiplying by \(g\) leaves the \((d{+}1)\)-st Gowers uniformity norm unchanged: writing \(f \cdot g\) for the pointwise product \((f \cdot g)(x) = f(x)\, g(x)\),
For \(f : \{ 0,1\} ^n \to \mathbb {R}\), base point \(x \in \{ 0,1\} ^n\), and direction vectors \(h_1, \ldots , h_k \in \{ 0,1\} ^n\), the order-\(k\) Gowers product is defined recursively:
It averages \(f\) over the \(2^k\) vertices of the combinatorial cube \(x + \mathrm{span}\{ h_1,\ldots ,h_k\} \).
Let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions, fix an order \(k \ge 0\), a base point \(x \in \{ 0,1\} ^n\), and direction vectors \(h_1, \ldots , h_k \in \{ 0,1\} ^n\), and write \(f \cdot g\) for the pointwise product \((f \cdot g)(y) = f(y)\, g(y)\). Then the order-\(k\) Gowers product of the product function equals the product of the two Gowers products:
Let \(f, g : \{ 0,1\} ^n \to \mathbb {R}\) be Boolean functions, and suppose \(g\) has degree \(\le d\), meaning that for every base point \(y \in \{ 0,1\} ^n\) and all direction vectors \(h_1,\ldots ,h_{d+1} \in \{ 0,1\} ^n\) one has \(\mathrm{GP}(g, d{+}1, y, h_1,\ldots ,h_{d+1}) = 1\). Then for every base point \(x \in \{ 0,1\} ^n\) and all directions \(h_1,\ldots ,h_{d+1} \in \{ 0,1\} ^n\), the order-\((d{+}1)\) Gowers product of the pointwise product \(f \cdot g\) agrees with that of \(f\) alone:
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function, and let \(F : \{ 0,1\} ^n \to \mathbb {R}\), \(F(x) = (-1)^{f(x)}\), be its \(\pm 1\) lift. Then for every order \(k \ge 0\), every base point \(x \in \{ 0,1\} ^n\), and all direction vectors \(h_1, \ldots , h_k \in \{ 0,1\} ^n\), the order-\(k\) Gowers product satisfies
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function, let \(k \ge 0\), let \(x \in \{ 0,1\} ^n\) be a base point, and let \(h_1, \ldots , h_{k+1} \in \{ 0,1\} ^n\) be direction vectors. Then the order-\((k{+}1)\) Gowers product splits along its last direction as
where \(x \oplus h_{k+1}\) denotes the componentwise XOR.
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function, and write \(x \oplus y\) for the componentwise XOR of \(x, y \in \{ 0,1\} ^n\). Then the fourth power of the second Gowers uniformity norm of \(f\) equals the squared \(L^2\) norm of the self-convolution \(f * f\), where \((f * f)(x) = \mathbb {E}_{y}[f(y)\, f(x \oplus y)]\); that is,
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function, with Gowers uniformity norm \(\| f\| _{U^2}\) and Fourier coefficients \(\hat f(S) = \mathbb {E}_x[f(x)\chi _S(x)]\) indexed by the subsets \(S \subseteq [n]\). Then
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function, let \(d \ge 1\), and let \(\varepsilon \) satisfy \(0 \le \varepsilon \le 1\). Suppose \(f\) is \(\varepsilon \)-far from degree \(\le d\); that is, \(\operatorname{dist}(f,g) \ge \varepsilon \) for every Boolean function \(g\) of degree \(\le d\). Then the \(\pm 1\) lift \((-1)^f : \{ 0,1\} ^n \to \{ \pm 1\} \) satisfies
where \(\lVert \, \cdot \, \rVert _{U^2}^{4}\) denotes the second Gowers uniformity norm raised to the fourth power, \(\mathbb {E}_{x}\, \mathbb {E}_{h_1,h_2}\! \bigl[\mathrm{GP}\bigl((-1)^f,2,x,h_1,h_2\bigr)\bigr]\).
A function \(f : \{ 0,1\} ^n \to \mathbb {R}\) has degree \(\le d\) if all its \((d{+}1)\)-fold multiplicative derivatives vanish: for every \(x\) and every \(h_1,\ldots ,h_{d+1}\), \(\mathrm{GP}(f, d{+}1, x, h_1,\ldots ,h_{d+1}) = 1\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be linear, meaning \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), where \(\oplus \) denotes componentwise XOR. Then \(f\) has degree \(\le 1\): its \(\pm 1\) lift \((-1)^f\) has vanishing multiplicative derivatives of order two, so that for every base point \(x \in \{ 0,1\} ^n\) and every pair of directions \(h_1, h_2 \in \{ 0,1\} ^n\) the order-\(2\) Gowers product of \((-1)^f\) over the combinatorial square \(x + \operatorname {span}\{ h_1, h_2\} \) equals \(1\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a linear Boolean function, meaning that \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), where \(\oplus \) denotes componentwise XOR. Then \(f\) belongs to the first-order Reed–Muller code \(\mathrm{RM}(1, n)\); that is, \(f\) has degree at most \(1\).
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\) be a Boolean function, and let \(x, h \in \{ 0,1\} ^n\). Then the multiplicative derivative of \(f\) in direction \(h\), evaluated at \(x\), coincides with the order-\(1\) Gowers product of \(f\) at base point \(x\) with the single direction vector \(h\); that is,
where \(\oplus \) denotes componentwise XOR.
For any constant \(c \in \mathbb {R}\) and any \(k\), the uniform average over \(k\) points \(h_1,\ldots ,h_k\) of the Boolean hypercube \(\{ 0,1\} ^n\) of the constant function equal to \(c\) is itself \(c\):
For every subset \(S \subseteq [n]\) and every integer \(d \ge 1\), the negated Walsh character \(x \mapsto -\chi _S(x)\) has degree \(\le d\). Equivalently, all of its \((d{+}1)\)-fold multiplicative derivatives are trivial: for every base point \(x \in \{ 0,1\} ^n\) and all direction vectors \(h_1,\ldots ,h_{d+1} \in \{ 0,1\} ^n\), the order-\((d{+}1)\) Gowers product satisfies \(\mathrm{GP}(-\chi _S,\, d{+}1,\, x,\, h_1,\ldots ,h_{d+1}) = 1\).
For every subset \(S \subseteq [n]\), the negated Walsh–Fourier character \(x \mapsto -\chi _S(x)\) on \(\{ 0,1\} ^n\) has degree \(\le 1\): its second-order multiplicative derivatives all equal \(1\). Explicitly, for every base point \(x \in \{ 0,1\} ^n\) and every pair of direction vectors \(h_1, h_2 \in \{ 0,1\} ^n\), the order-\(2\) Gowers product of \(-\chi _S\) satisfies \(\mathrm{GP}(-\chi _S, 2, x, h_1, h_2) = 1\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function, let \(d \ge 1\), and write \(F = (-1)^f : \{ 0,1\} ^n \to \{ \pm 1\} \) for its \(\pm 1\) lift. Suppose \(f\) is \(\varepsilon \)-far from degree \(\le d\); that is, \(0 \le \varepsilon \le 1\) and every Boolean function \(g\) of degree \(\le d\) satisfies \(\Pr _{x}[f(x) \ne g(x)] \ge \varepsilon \). Then for every \(S \subseteq [n]\),
where \(\widehat{F}(S) = \mathbb {E}_x\bigl[F(x)\, \chi _S(x)\bigr]\) is the Fourier coefficient of \(F\) at \(S\); equivalently, \(\widehat{F}(S) \ge 2\varepsilon - 1\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function, and let \(g : \{ 0,1\} ^n \to \mathbb {R}\) be its \(\pm 1\) lift \(g(x) = (-1)^{f(x)}\). Writing \(\hat g(S) = \mathbb {E}[g\, \chi _S]\) for the Fourier coefficient of \(g\) at \(S\), where \(\chi _S(x) = \prod _{i \in S}(-1)^{x_i}\) and the expectation is over \(x\) uniform on \(\{ 0,1\} ^n\), one has
Fix an ambient dimension \(n\) and a degree \(d\). Every Boolean function \(f : \{ 0,1\} ^n \to \{ 0,1\} \) of degree at most \(d\) also has degree at most \(d+1\); equivalently, the Reed–Muller code \(\mathrm{RM}(d,n)\) is contained in \(\mathrm{RM}(d+1,n)\).
Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function lying in the Reed–Muller code \(\mathrm{RM}(1,n)\), that is, a function of degree at most \(1\). Then \(f\) is affine: either \(f\) is \(\mathbb {F}_2\)-linear, meaning \(f(x \oplus y) = f(x) \oplus f(y)\) for all \(x, y \in \{ 0,1\} ^n\), or the complementary function \(x \mapsto \lnot f(x)\) is \(\mathbb {F}_2\)-linear.
Fix a degree bound \(d\) and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function belonging to the Reed–Muller code \(\mathrm{RM}(d,n)\), that is, a function of degree at most \(d\). Then the \((d+1)\)-fold derivative test accepts \(f\) with probability exactly \(1\):
where \(x\) and the direction vectors \(h_1,\ldots ,h_{d+1}\) range uniformly over \(\{ 0,1\} ^n\).
Fix \(d \ge 1\) and let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function. Suppose \(f\) is \(\varepsilon \)-far from degree \(\le d\); that is, \(0 \le \varepsilon \le 1\) and \(\operatorname{dist}(f,g) \ge \varepsilon \) for every Boolean function \(g\) of degree \(\le d\). Then the \((d{+}1)\)-fold derivative test accepts \(f\) with probability at most \(1 - \varepsilon \).
For every \(x \in \{ 0,1\} ^n\), the componentwise XOR of \(x\) with the all-zeros vector \(\mathbf{0}\) equals \(x\); that is, \(x \oplus \mathbf{0} = x\). Equivalently, \(\mathbf{0}\) is a right identity for the group operation on \((\mathbb {F}_2)^n\).
Let \(k \ge 1\) and \(n \ge 2\) be integers, and let \(\delta , \varepsilon \in (0,1]\). Suppose \(L\) is a normal form \((3,\delta ,\varepsilon )\)-decodable code with message length \(k\) and block length \(n\), whose density parameter equals \(\delta \) and whose advantage parameter equals \(\varepsilon \). Then there exists a real constant \(C \gt 0\) for which
Here \(C\) is quantified after \(k\), \(n\), \(\delta \), \(\varepsilon \) and the code \(L\) have been fixed, so it is permitted to depend on all of them.
Let \(X\colon \Omega \to \mathbb {R}\) be a measurable random variable on a measure space \((\Omega , \mu )\) whose law is the centered Gaussian \(N(0, \sigma _0)\) with variance \(\sigma _0 \ge 0\). Then for every scalar \(c \in \mathbb {R}\), the scaled variable \(c\, X\) has law
the centered Gaussian with variance \(c^2 \sigma _0\).
B&M’s unimodular (Thm 12.3, p. 226): every full square submatrix (order \(\nu \)\(-\)1, selected by an injective column map) has determinant 0, +1 or -1.
A matrix is said to be unimodular if all its full square submatrices have determinants 0, +1 or -1.
Let \(X\) be any index set and \(Y\) a finite index set, and let \(A, B \in \mathbb {R}^{X \times Y}\) be real matrices with rows indexed by \(X\) and columns indexed by \(Y\). Then
Let \(s\) be a finite subset of an index set, and let \((A_i)_{i \in s}\) be a family of real matrices with rows indexed by a set \(X\) and columns indexed by a finite set \(Y\). Then the rank of their sum is at most the sum of their ranks:
Let \(k\) be a positive integer, let \(\sigma ^2 \gt 0\) be a real number, and let \(d_1\) and \(d_2\) be positive integers. Then there exists a positive real number \(C_{\mathrm{MK}} \gt 0\) satisfying
Equip the sample space \(\Omega = \{ \omega : \mathrm{Fin}\, k \to \mathrm{Fin}\, d \to \mathbb {R}\} \) of \(k \times d\) real arrays with its product \(\sigma \)-algebra, and the space \(\mathbb {R}^{k \times d}\) of matrices likewise. Then the map \(A : \Omega \to \mathbb {R}^{k\times d}\) sending \(\omega \) to the matrix with entries \(A(\omega )_{ij} = \omega _{ij}/\sqrt{k}\) is measurable.
Fix a vector \(x\in \mathbb {R}^d\) and a real number \(\varepsilon \), and let \((\Omega ,\mathcal{F})\) be a measurable space. Suppose \(A\colon \Omega \to \mathbb {R}^{k\times d}\) is a measurable map assigning to each \(\omega \) a \(k\times d\) real matrix \(A(\omega )\). Then the set of outcomes on which \(x\) is badly distorted,
is measurable.
Let \(a\), \(b\), \(c\) be elements of a linearly ordered set. Then the value returned by the timed computation \(\mathrm{median3}(a,b,c)\) is a median of \(a\), \(b\), and \(c\): it is one of the three inputs occupying the middle position in the order, so that at least one of the remaining two values is \(\le \) it and at least one is \(\ge \) it.
Let \(a\), \(b\), \(c\) be elements of a linear order, and let \(\mathrm{median3}(a,b,c)\) be the timed computation that returns a median of the three by a nested sequence of comparisons, charging one unit of time for each comparison performed. Then the recorded time cost of \(\mathrm{median3}(a,b,c)\) is at most \(3\); that is, the computation performs at most three comparisons.
Fix \(k \gt 0\) and a vector \(x \in \mathbb {R}^d\), and let \(\mu \) be the joint Rademacher product measure on the sample space of \(k \times d\) arrays \(\omega = (\omega _{ij})\). Then for every pair of indices \(i\) and \(j\), the coordinate function \(\omega \mapsto \dfrac {x_j}{\sqrt{k}}\, \omega _{ij}\) is square-integrable with respect to \(\mu \); that is, it belongs to \(L^2(\mu )\).
Let \(Z\) be a measurable real random variable on a probability space \((\Omega , \mu )\) which is strictly sub-Gaussian with variance proxy \(\sigma ^2 \gt 0\) (for every real \(s\), \(e^{sZ}\) is integrable and \(\mathbb {E}_\mu [e^{sZ}] \le e^{\sigma ^2 s^2/2}\)). Then for every \(t\) with \(0 \le t \le 1/(4\sigma ^2)\), the function \(e^{t(Z^2 - \sigma ^2)}\) is \(\mu \)-integrable and
Let \(Z\) be a measurable real random variable on a probability space \((\Omega , \mu )\) which is strictly sub-Gaussian with variance proxy \(\sigma ^2 \gt 0\) (for every real \(s\), \(e^{sZ}\) is integrable and \(\mathbb {E}_\mu [e^{sZ}] \le e^{\sigma ^2 s^2/2}\)), with \(\mathbb {E}_\mu [Z^2] = \sigma ^2\). Then for every \(t\) with \(-1/(4\sigma ^2) \le t \le 0\), the function \(e^{t(Z^2 - \sigma ^2)}\) is \(\mu \)-integrable and
Let \(G\) be a multigraph, and let \(C\) be a minimum cut of \(G\), given by a vertex subset \(S\) that is nonempty and has nonempty complement in \(V(G)\). Suppose \(e \in E(G)\) is an edge that does not cross \(C\), so that both endpoints of \(e\) lie on the same side of the partition. Then the cut induced by \(C\) on the contracted multigraph \(G/e\) — obtained by merging the endpoints of \(e\) and intersecting \(S\) with the surviving vertices — is again a minimum cut, of \(G/e\).
Let \(G\) be a finite two-player zero-sum game with payoff matrix \(A\), having at least two row actions and at least one column action, and suppose every entry satisfies \(0 \le A_{ij} \le 1\). Then for every \(\varepsilon \gt 0\) there exist a mixed row strategy \(p\) and a mixed column strategy \(q\) such that, for every pure row \(i\) and every pure column \(j\),
That is, the row player’s guaranteed payoff under \(p\) against any column falls short of the column player’s exposure under \(q\) against any row by at most \(\varepsilon \).
A mixed strategy over \(n\) actions is a record consisting of a weight function \(w : \mathrm{Fin}\, n \to \mathbb {R}\) together with proofs that every weight is non-negative and that the weights sum to \(1\), i.e. \(w\) is a probability distribution over \(\mathrm{Fin}\, n\).
Let \(D\) be a diconnected tournament on a finite vertex set with at least three vertices, and let \(u\) be any vertex of \(D\). Then for every integer \(k\) with \(3 \le k \le \left\lvert V(D)\right\rvert \) there is a directed cycle of length \(k\) passing through \(u\).
For every real \(\delta \) with \(0 \le \delta \le 1/2\), the asymptotic binary code rate satisfies
where \(H\) is the base-\(2\) binary entropy function.
If \(e \in E(G)\) does not cross the cut \(C\), then \(C\) induces a cut of the contracted multigraph \(G/e\): the vertex side is the image of \(S\) after merging the endpoints of \(e\), and both nonemptiness conditions continue to hold because \(e\) has both endpoints on the same side.
The multiset of edges of \(G\) that cross the cut \(C\), i.e. those edges having exactly one endpoint in \(S\), counted with multiplicity. It is obtained by filtering the edge multiset by the condition that exactly one of the two representative endpoints lies in \(S\).
A multigraph on a vertex type \(\alpha \) consists of a finite vertex set together with a multiset of unordered pairs (elements of \(\mathrm{Sym}_2\, \alpha \)) as edges, subject to two conditions: every endpoint of every edge is a live vertex, and no edge is a self-loop (edges are loopless, since loops are deleted after each contraction).
Given an edge \(e \in E(G)\) with endpoints \(u\) and \(v\), the contracted multigraph \(G/e\) is obtained by replacing every occurrence of \(v\) by \(u\) in all edges, deleting the resulting self-loops (in particular all parallel copies of \(e\)), and removing \(v\) from the vertex set. Parallel edges between the merged vertex and the rest of the graph are retained with multiplicity.
Given three Boolean values \(a, b, c\) (the truth values of a clause’s three literals) and an index \(k \in \mathrm{Fin}\, 3\), returns the color in \(\mathrm{Fin}\, 3\) assigned to the \(k\)-th internal clause-gadget node. The coloring is chosen so that all three gadget nodes receive distinct colors whenever \((a, b, c)\) is a NAE-satisfying assignment; if the assignment is not NAE-satisfying, all gadget nodes are colored \(0\).
Given a NAE-SAT instance \(\mathit{clauses}\), the directed edge relation on \(\texttt{NAEtoColor.OutputVertex}\; V\) is defined by: (1) \(\mathtt{groundNode}\) is adjacent to every \(\mathtt{varNode}\); (2) \(\mathtt{varNode}\; v\) is adjacent to \(\mathtt{clauseNode}\; c\; i\) when \(v\) is the \(i\)-th variable of clause \(c\); (3) two clause-gadget nodes \(\mathtt{clauseNode}\; c\; i\) and \(\mathtt{clauseNode}\; c\; j\) are adjacent when \(c \in \mathit{clauses}\) and \(i \neq j\) (forming a triangle within each clause gadget). All other pairs are non-adjacent.
A Not-All-Equal clause over a variable type \(V\) is an ordered triple \((v_0, v_1, v_2)\) of variables from \(V\). It records the three variables of one NAE constraint; such a clause is meant to be satisfied by a Boolean assignment exactly when the values assigned to \(v_0\), \(v_1\), \(v_2\) are not all equal.
Given a Boolean assignment \(\mathit{assign} : V \to \mathtt{Bool}\), defines a coloring of all vertices of the reduction graph by: \(\mathtt{groundNode} \mapsto 0\); \(\mathtt{varNode}\; v \mapsto 1\) if \(\mathit{assign}(v) = \mathtt{true}\), else \(2\); \(\mathtt{clauseNode}\; c\; k \mapsto \texttt{NAEtoColor.clauseNodeColor}\; (\mathit{assign}(c.v_0))\; (\mathit{assign}(c.v_1))\; (\mathit{assign}(c.v_2))\; k\).
Let \(f\) be a NAE-SAT instance over a variable type \(V\), that is, a list of not-all-equal clauses each naming three variables. If \(f\) is satisfiable — some Boolean assignment \(V \to \{ \mathrm{true}, \mathrm{false}\} \) makes the three variables of every clause not all receive the same value — then the reduction graph of \(f\) is 3-colorable, meaning it admits a proper coloring with the three colors \(\{ 0, 1, 2\} \).
Let \(f\) be a NAE-SAT instance over a variable type \(V\), and let \(\Gamma (f)\) denote its reduction graph, the simple graph built from \(f\) as described above (a ground vertex joined to every variable vertex, each variable vertex joined to the clause-gadget vertices for the clauses in which it occurs, and the three gadget vertices of each clause joined pairwise into a triangle). Then \(f\) is satisfiable if and only if \(\Gamma (f)\) is 3-colorable.
Let \(f\) be a NAE-SAT instance over a variable type \(V\), that is, a list of Not-All-Equal clauses, each clause naming three variables. Form the reduction graph associated to \(f\): its vertices are a single ground vertex, one vertex for each variable in \(V\), and, for every clause, three gadget vertices; its edges join the ground vertex to each variable vertex, join each variable vertex \(v\) to a clause-gadget vertex whenever \(v\) occurs in the corresponding position of that clause, and, for every clause of \(f\), join its three gadget vertices pairwise into a triangle. If this graph is 3-colorable — that is, its vertices admit a coloring by \(\{ 0,1,2\} \) in which adjacent vertices receive distinct colors — then \(f\) is satisfiable: there exists a Boolean assignment \(\mathit{assign} : V \to \{ \mathrm{true},\mathrm{false}\} \) under which every clause of \(f\) has its three assigned values not all equal.
The vertex set of the reduction graph is an inductive type with three constructors:
\(\mathtt{groundNode}\) — a single ground vertex (colored “neutral”);
\(\mathtt{varNode}\; v\) — one vertex per variable \(v \in V\);
\(\mathtt{clauseNode}\; c\; k\) — three internal gadget vertices (\(k \in \mathrm{Fin}\, 3\)) for each clause \(c\).
Given a NAE-SAT instance \(f\), \(\texttt{NAEtoColor.ReductionGraph}\; f\) is the simple graph on \(\texttt{NAEtoColor.OutputVertex}\; V\) whose adjacency is the symmetrization of \(\texttt{NAEtoColor.EdgeRelation}\; f\) restricted to distinct pairs (no self-loops).
Given a Boolean assignment \(\mathit{assign} : V \to \mathtt{Bool}\) and a Not-All-Equal clause \(c\) with three variables \(c.v_0, c.v_1, c.v_2\), \(\texttt{NAEtoColor.SatisfiesClause}\) returns \(\mathtt{true}\) if and only if the three assigned values are not all equal, i.e. at least two of \(\mathit{assign}(c.v_0)\), \(\mathit{assign}(c.v_1)\), \(\mathit{assign}(c.v_2)\) differ.
Given an assignment \(\mathit{assign} : V \to \mathtt{Bool}\) and a NAE-SAT instance \(f\), \(\texttt{NAEtoColor.SatisfiesNAE3}\) returns \(\mathtt{true}\) if and only if every clause in \(f\) is satisfied by \(\mathit{assign}\) (i.e. \(\texttt{NAEtoColor.SatisfiesClause}\) holds for each clause).
The associated digraph D(G): each edge becomes two opposite arcs (cited as ex 10.3.6).
The associated digraph D(G) of a graph G is the digraph obtained when each edge e of G is replaced by two oppositely oriented arcs with the same ends as e.
Let \(G\) be a finite simple graph, and let \(k \ge 2\) be an integer such that \(G\) is \(k\)-connected, that is, \(k \le \kappa (G)\). Then for any set \(T\) of exactly \(k\) vertices of \(G\), there is a cycle in \(G\) whose vertices include every vertex of \(T\).
Let \(N\) be a network with a single source \(x\) and a single sink \(y\), and let \(f\) be a flow in \(N\). Suppose \(N\) contains no \(f\)-incrementing path from \(x\) to \(y\) — that is, no path from the source to the sink each of whose arcs is either a forward arc that is not yet saturated (\(f \lt c\)) or a backward arc carrying positive flow (\(f \gt 0\)). Then there is a cut \((S,\bar S)\), a set of vertices with \(x \in S\) and \(y \notin S\), whose capacity equals the value of \(f\):
Let \(N\) be a network with distinguished single source \(x\) and single sink \(y\). Then \(N\) possesses a minimum cut: there exists a set \(S\) of vertices with \(x \in S\) and \(y \notin S\) whose capacity \(\sum _{u \in S}\sum _{v \notin S} c(u,v)\) is least among all sets separating \(x\) from \(y\) in this way.
Let \(N\) be a network with disjoint nonempty source set \(X\) and sink set \(Y\), and let \(f\) be a flow in \(N\). For a vertex set \(S\), write \(f^{+}(S)\) for the total flow along the arcs directed from \(S\) to its complement and \(f^{-}(S)\) for the total flow along the arcs directed from the complement into \(S\). Then
(Gale, 1957.) Let \(N\) be a network with source set \(X\), sink set \(Y\), and capacity function \(c\), and suppose each vertex \(x\) is assigned a non-negative integer supply \(\sigma (x)\) and each vertex \(y\) a non-negative integer demand \(\partial (y)\). Then \(N\) admits a feasible flow for the supplies \(\sigma \) and demands \(\partial \) if and only if, for every vertex set \(S\),
where \(\bar S\) denotes the complement of \(S\) and \(\mathrm{cap}(S,\bar S)\) is the capacity of the cut from \(S\) to \(\bar S\).
Let \(p=(p_1,\dots ,p_m)\) and \(q=(q_1,\dots ,q_n)\) be sequences of non-negative integers with equal total, \(\sum _{i=1}^{m} p_i = \sum _{j=1}^{n} q_j\), and suppose \(q\) is non-increasing, \(q_1 \ge q_2 \ge \cdots \ge q_n\). Then \((p,q)\) is realisable by a simple bipartite graph if and only if
Let \(N\) be a network with single source \(x\) and single sink \(y\), let \(f\) be a flow in \(N\), and let \((S, \bar S)\) be a cut of \(N\), so that \(x \in S\) and \(y \notin S\). If the value of the flow equals the capacity of the cut, \(\mathrm{val}(f) = \mathrm{cap}(S)\), then \(f\) is a maximum flow in \(N\) and \((S,\bar S)\) is a minimum cut in \(N\).
Let \(N\) be a network with a single source \(x\) and a single sink \(y\), and recall that a cut is determined by a vertex set \(S\) with \(x \in S\) and \(y \notin S\), its capacity being \(\sum _{u \in S}\sum _{v \notin S} c(u,v)\); a cut is a minimum cut if no cut has strictly smaller capacity. If the cuts determined by vertex sets \(S\) and \(T\) are both minimum cuts of \(N\), then the cuts determined by \(S \cup T\) and by \(S \cap T\) are also minimum cuts of \(N\).
G is (m,n)-orientable: it admits an orientation in which every indegree is m or n.
(B&M exercise 11.5.5). An (m+n)-regular graph G is (m,n)-orientable if it can be oriented so that each indegree is either m or n.
Let \(G\) be a finite simple graph that is \((m+n)\)-regular. Then \(G\) is \((m,n)\)-orientable — it admits an orientation in which every vertex has indegree \(m\) or \(n\) — if and only if the vertex set can be split into two disjoint sets \(V_1\) and \(V_2\) whose union is the whole vertex set, such that for every set \(S\) of vertices
where \(\left\lvert [S,\bar S]\right\rvert \) is the number of edges of \(G\) with one end in \(S\) and the other outside \(S\).
Let \(G\) be a finite \((m+n)\)-regular simple graph, and suppose \(n \lt m\). If \(G\) is \((m,n)\)-orientable — that is, its edges admit an orientation in which every vertex has indegree equal to \(m\) or to \(n\) — then \(G\) is also \((m-1,\, n+1)\)-orientable.
Let \(G\) be a simple graph on a finite vertex set \(V\), and let \(k\) be a natural number with \(|V| \ge k+1\). Then \(\kappa (G) \ge k\) if and only if every pair of distinct vertices \(u, v\) of \(G\) is joined by at least \(k\) internally disjoint \(u\)–\(v\) paths.
Let \(G\) be a finite simple graph and let \(k\) be a natural number. Then the edge connectivity satisfies \(k \le \kappa '(G)\) if and only if, for every pair of distinct vertices \(u\) and \(v\), the maximum number of pairwise edge-disjoint paths joining \(u\) and \(v\) is at least \(k\).
Let \(N\) be a network with distinguished single source \(x\) and single sink \(y\), and suppose there is no directed path from \(x\) to \(y\) every arc of which has positive capacity. Then every maximum flow \(f\) in \(N\) has value \(0\), and every minimum cut \(S\) (a vertex set with \(x \in S\) and \(y \notin S\)) has capacity \(0\).
Let \(N\) be a network — a finite digraph carrying a nonnegative integer capacity on each arc, with disjoint nonempty source and sink sets and a distinguished source \(x\) and sink \(y\) — and let \(f\) be a flow in \(N\). Then \(f\) is a maximum flow if and only if there is no \(f\)-incrementing path in \(N\) ending at the sink \(y\).
Let \(N\) be a network on a finite vertex set, with capacity function \(c\), disjoint nonempty source and sink sets \(X\) and \(Y\), and a distinguished source \(x\) and sink \(y\) (Ford and Fulkerson, 1956). Then \(N\) admits a maximum flow \(f\) together with a minimum cut \((S,\bar S)\) — that is, \(x \in S\) and \(y \notin S\) — whose value and capacity coincide:
In particular, the value of a maximum flow in \(N\) equals the capacity of a minimum cut.
Let \(D\) be a finite digraph, and let \(x\) and \(y\) be two of its vertices. Then the maximum number of pairwise arc-disjoint directed paths from \(x\) to \(y\) equals the minimum number of arcs whose deletion destroys every directed path from \(x\) to \(y\).
Let \(G\) be a finite simple graph and let \(x\) and \(y\) be two of its vertices. Then the maximum number of pairwise edge-disjoint \((x,y)\)-paths in \(G\) equals the minimum number of edges whose deletion destroys every \((x,y)\)-path in \(G\).
Let \(D\) be a finite digraph and let \(x\) and \(y\) be two vertices of \(D\) with no arc from \(x\) to \(y\). Then the maximum number of pairwise internally-disjoint directed paths from \(x\) to \(y\) equals the minimum number of vertices (other than \(x\) and \(y\)) whose deletion destroys every directed path from \(x\) to \(y\).
Let \(G\) be a finite simple graph, and let \(x\) and \(y\) be two distinct, nonadjacent vertices of \(G\). Then the maximum number of internally disjoint \((x,y)\)-paths in \(G\) equals the minimum size of a set of vertices (necessarily distinct from \(x\) and \(y\)) whose deletion destroys every \((x,y)\)-path.
A network N = (D, X, Y, c).
A network N is a digraph D (the underlying digraph of N) with two distinguished subsets of vertices, X and Y, and a non-negative integer-valued function c defined on its arc set A; the sets X and Y are assumed to be disjoint and nonempty. The vertices in X are the sources of N and those in Y are the sinks of N. They correspond to production centres and markets, respectively. … The function c is the capacity function of N and its value on an arc a the capacity of a. The capacity of an arc can be thought of as representing the maximum rate at which a commodity can be transported along it.
cap (S, \(\bar{S}\)) = \(\sum \)_{u\(\in \)S} \(\sum \)_{v\(\notin \)S} c u v.
The capacity of a cut K is the sum of the capacities of its arcs. We denote the capacity of K by cap K; thus cap K = \(\sum \)_{a \(\in \) K} c(a).
f\(^{-}\)(S) = f(\(\bar{S}\), S) = \(\sum \)_{u\(\notin \)S} \(\sum \)_{v\(\in \)S} f u v.
…and f\(^{-}\)(S) for f(\(\bar{S}\), S).
The book names the difference: If S is a subset of vertices in a network N and f is a flow in N, then f\(^{+}\)(S) - f\(^{-}\)(S) is called the resultant flow out of S, and f\(^{-}\)(S) - f\(^{+}\)(S) the resultant flow into S, relative to f.
f\(^{+}\)(S) = f(S, \(\bar{S}\)) = \(\sum \)_{u\(\in \)S} \(\sum \)_{v\(\notin \)S} f u v.
If S \(\subseteq \) V, we denote V \ S by \(\bar{S}\). … If f is a real-valued function defined on the arc set A of N, and if K \(\subseteq \) A, we denote \(\sum \)_{a \(\in \) K} f(a) by f(K). Furthermore, if K is a set of arcs of the form (S, \(\bar{S}\)), we shall write f\(^{+}\)(S) for f(S, \(\bar{S}\)) and f\(^{-}\)(S) for f(\(\bar{S}\), S).
An f-incrementing path from the source.
The path P is f-unsaturated if \(\iota \)(P) > 0 (or, equivalently, if each forward arc of P is f-unsaturated and each reverse arc of P is f-positive). Put simply, an f-unsaturated path is one that is not being used to its full capacity. An f-incrementing path is an f-unsaturated path from the source x to the sink y.
The residual capacity \(\iota \)(P) > 0 of an incrementing path — a recursion over IncPath.
With each path P in N we associate a non-negative integer \(\iota \)(P) = min_{a \(\in \) A(P)} \(\iota \)(a), where \(\iota \)(a) = c(a) - f(a) if a is a forward arc of P, and \(\iota \)(a) = f(a) if a is a reverse arc. As may easily be seen, \(\iota \)(P) is the largest amount by which the flow along P can be increased (relative to f) without violating condition (11.1).
A cut (S, \(\bar{S}\)) with x \(\in \) S, y \(\notin \) S.
Let N be a network with a single source x and a single sink y. A cut in N is a set of arcs of the form (S, \(\bar{S}\)), where x \(\in \) S and y \(\in \) \(\bar{S}\).
Feasibility for supplies \(\sigma \) and demands dem (B&M’s \(\partial \)).
Suppose that to each source x\(_{i}\) of N is assigned a non-negative integer \(\sigma \)(x\(_{i}\)), called the supply at x\(_{i}\), and to each sink y\(_{j}\) of N is assigned a non-negative integer \(\partial \)(y\(_{j}\)), called the demand at y\(_{j}\). A flow f in N is said to be feasible if
f\(^{+}\)(x\(_{i}\)) - f\(^{-}\)(x\(_{i}\)) \(\le \) \(\sigma \)(x\(_{i}\)) for all x\(_{i}\) \(\in \) X f\(^{-}\)(y\(_{j}\)) - f\(^{+}\)(y\(_{j}\)) \(\ge \) \(\partial \)(y\(_{j}\)) for all y\(_{j}\) \(\in \) Y
In other words, a flow f is feasible if the resultant flow out of each source x\(_{i}\) relative to f does not exceed the supply at x\(_{i}\), and the resultant flow into each sink y\(_{j}\) relative to f is at least as large as the demand at y\(_{j}\).
A flow: (11.1) capacity constraint + (11.2) conservation at intermediate vertices.
A flow in a network N is an integer-valued function f defined on A such that
0 \(\le \) f(a) \(\le \) c(a) for all a \(\in \) A (11.1) f\(^{-}\)(v) = f\(^{+}\)(v) for all v \(\in \) I (11.2)
The upper bound in condition (11.1) is called the capacity constraint; it imposes the natural restriction that the rate of flow along an arc cannot exceed the capacity of the arc. Condition (11.2), called the conservation condition, requires that, for any intermediate vertex v, the rate at which material is transported into v is equal to the rate at which it is transported out of v. Note that every network has at least one flow, since the function f defined by f(a) = 0, for all a \(\in \) A, clearly satisfies both (11.1) and (11.2); it is called the zero flow.
Min number of arcs whose deletion destroys all directed (x,y)-paths in a Network.
(B&M lemma 11.4(b), verbatim). …the minimum number n of arcs whose deletion destroys all directed (x, y)-paths in N.
The revised flow \(\hat{f}\) (11.9) obtained by pushing \(\iota \)(P) along P.
By sending an additional flow of \(\iota \)(P) along P, one obtains a new flow \(\hat{f}\) defined by \(\hat{f}\)(a) = f(a) + \(\iota \)(P) if a is a forward arc of P, \(\hat{f}\)(a) = f(a) - \(\iota \)(P) if a is a reverse arc of P, and \(\hat{f}\)(a) = f(a) otherwise.
val f = f\(^{+}\)(X) \(-\) f\(^{-}\)(X).
Since the conservation condition requires that the resultant flow out of any intermediate vertex is zero, it is intuitively clear and not difficult to show (exercise 11.1.3) that, relative to any flow f, the resultant flow out of X is equal to the resultant flow into Y. This common quantity is called the value of f, and is denoted by val f; thus
val f = f\(^{+}\)(X) - f\(^{-}\)(X)
A flow f in N is a maximum flow if there is no flow f’ in N such that val f’ > val f.
Let \(p_0, p_1, \ldots , p_m\) and \(q_0, q_1, \ldots , q_{n-1}\) be nonnegative integers, regarded as prescribed row degrees and column degrees. Then the pair \((p, q)\) is realisable by a simple bipartite graph if and only if the reduced pair \((p', q')\) is realisable, where \(p' = (p_1, \ldots , p_m)\) is obtained by deleting the first row degree \(p_0\), and the reduced column degrees are
that is, one is subtracted (never below zero) from each of the first \(p_0\) column degrees and the remaining column degrees are left unchanged.
(p,q) is realisable by a simple bipartite graph: there is a (0,1)-matrix B with row sums p and column sums q.
We say that the pair (p, q) is realisable by a simple bipartite graph if there exists a simple bipartite graph G with bipartition ({x\(_{1}\), x\(_{2}\), …, x_m}, {y\(_{1}\), y\(_{2}\), …, y_n}), such that d(x\(_{i}\)) = p\(_{i}\) for 1 \(\le \) i \(\le \) m and d(y\(_{j}\)) = q\(_{j}\) for 1 \(\le \) j \(\le \) n.
The vertex-splitting digraph D’ of Theorem 11.6, on V \(\oplus \) V: inl v = v’, inr v = v”, with internal arcs v’ \(\to \) v” and each original arc u \(\to \) v redrawn as u” \(\to \) v’.
Construct a new digraph D’ from D as follows: (i) split each vertex v \(\in \) V \ {x, y} into two new vertices v’ and v”, and join them by an arc (v’, v”); (ii) replace each arc of D with head v \(\in \) V \ {x, y} by a new arc with head v’, and each arc of D with tail v \(\in \) V \ {x, y} by a new arc with tail v”.
Let \(N\) be a network on a finite vertex set \(V\), let \(f\) assign a nonnegative integer \(f(u,v)\) to each ordered pair of vertices \(u,v\), and let \(S\subseteq V\). Summing the resultant flow out of a singleton \(\{ v\} \) over all vertices \(v\in S\) recovers the resultant flow out of \(S\); that is, the following identity holds in \(\mathbb {Z}\):
Let \(N\) be a network with distinguished source \(x\) and sink \(y\) whose capacity function takes only the values \(0\) and \(1\) (that is, every capacity is at most \(1\)), and let \(f\) be a maximum flow in \(N\). Then the value of \(f\) equals the maximum number of arc-disjoint directed \((x,y)\)-paths in \(N\).
Let \(N\) be a network with capacity function \(c\), a single source \(x\), and a single sink \(y\), and suppose every arc has capacity at most \(1\), that is \(c(u,v) \le 1\) for all vertices \(u,v\). For a set \(S\) of vertices with \(x \in S\) and \(y \notin S\), write \(\mathrm{cap}(S) = \sum _{u \in S}\sum _{v \notin S} c(u,v)\) for its cut capacity, and suppose \(S\) is a minimum cut, meaning \(\mathrm{cap}(S) \le \mathrm{cap}(S')\) for every set \(S'\) with \(x \in S'\) and \(y \notin S'\). Then \(\mathrm{cap}(S)\) equals the minimum number of arcs whose deletion destroys every directed \((x,y)\)-path in \(N\).
Let \(N\) be a network with capacity function \(c\), distinguished source \(x\) and sink \(y\), let \(f\) be a flow in \(N\), and let \(S\) be a vertex set defining a cut, so that \(x \in S\) and \(y \notin S\); write \(\bar S\) for the complementary vertex set. Then the value of the flow equals the capacity of the cut, \(\mathrm{val}(f) = \mathrm{cap}(S,\bar S)\), if and only if \(f(u,v) = c(u,v)\) for every arc with \(u \in S\) and \(v \in \bar S\) (each forward arc across the cut is saturated) and \(f(u,v) = 0\) for every arc with \(u \in \bar S\) and \(v \in S\) (each backward arc across the cut carries no flow).
Let \(N\) be a network with source set \(X\), a single distinguished source vertex \(x\), and a single distinguished sink vertex \(y\); let \(f\) be a flow in \(N\); and let \(S\) be a set of vertices forming a cut, so that \(x \in S\) and \(y \notin S\). Then the value \(\operatorname {val} f = f^{+}(X) - f^{-}(X)\) of the flow is equal to the resultant flow out of \(S\):
Let \(N\) be a network with disjoint nonempty source set \(X\) and sink set \(Y\), distinguished source \(x\) and sink \(y\), and integer capacity function \(c\). Let \(f\) be a flow in \(N\), and let \((S,\bar S)\) be a cut, so that \(x\in S\) and \(y\notin S\). Then the value of \(f\) is at most the capacity of the cut:
If G has at least one pair of distinct nonadjacent vertices, the connectivity \(\kappa \)(G) of G is the minimum k for which G has a k-vertex cut; otherwise, we define \(\kappa \)(G) to be \(\nu \) - 1. … G is said to be k-connected if \(\kappa \)(G) \(\ge \) k.
Let \(A\) be a \(k \times d\) real matrix, regarded as the linear map \(x \mapsto Ax\) from the Euclidean space \(\mathbb {R}^d\) to \(\mathbb {R}^k\), and let \(x \in \mathbb {R}^d\). Then the squared Euclidean norm of the image is the sum of the squares of its coordinates,
Let \(q \ge 2\) be an integer, let \(k\) and \(n\) be positive integers, and let \(\delta , \varepsilon \gt 0\). Then there exist a natural number \(n' \le 6n\) and real numbers \(\delta ', \varepsilon ' \gt 0\) such that
The structural data of a \((3,\delta ,\varepsilon )\)-normally decodable code with message length \(k\) and block length \(n\): for each message index \(i : \mathrm{Fin}\, k\) a hypergraph \(H_i\) on \(\mathrm{Fin}\, n\) which is \(3\)-uniform and a matching, together with real parameters \(\delta \gt 0\) and \(\varepsilon \gt 0\) and the density requirement \(\delta n \le \left\lvert H_i\right\rvert \) for every \(i\).
For a normal form code \(L\), a message \(b : \mathrm{Fin}\, k \to \mathbb {Z}\) and an assignment \(x : \mathrm{Fin}\, n \to \mathbb {Z}\), the real-valued XOR instance
where \(m\) is the total number of constraints of \(L\).
Let \(G\) be a multigraph with \(n\) vertices, and let \(\alpha \) and \(m\) be natural numbers with \(\alpha \ge 1\) and \(2\alpha \le n\). If, as an inequality of real numbers, \(m \cdot \binom {n}{2\alpha }^{-1} \le 1\), then \(m \le n^{2\alpha }\).
Given \(n \geq 1\), OnlineLearning.arbitraryMixedStrategy is a concrete element of \(\mathrm{MixedStrategy}(n)\) that places all probability mass on a single arbitrarily chosen action. It is used solely to witness that the type of mixed strategies is nonempty.
Let \(\iota \) be nonempty and let \(g : \iota \to \kappa \to \mathbb {R}\) take values in \([0,1]\), i.e. \(0 \le g(i,j) \le 1\) for all \(i, j\). Then the family \(j \mapsto \inf _{i \in \iota } g(i, j)\) has range bounded above.
Let \(\kappa \) be nonempty and let \(g : \iota \to \kappa \to \mathbb {R}\) take values in \([0,1]\), i.e. \(0 \le g(i,j) \le 1\) for all \(i, j\). Then the family \(i \mapsto \sup _{j \in \kappa } g(i, j)\) has range bounded below.
Let \(S \subseteq \mathbb {R}\) and let \(g : \mathbb {R}\to \mathbb {R}\) be concave on \(S\). Let \(p = (p_1, \dots , p_n)\) be a mixed strategy over \(n\) actions, that is, a tuple of weights \(p_j \ge 0\) with \(\sum _{j=1}^{n} p_j = 1\), and let \(y_1, \dots , y_n\) be points of \(S\). Then
Let \(\iota \) be a finite index set and let \(L \colon \mathbb {R}^{\iota } \to \mathbb {R}\) be a continuous linear functional. Writing \(e_i\) for the standard basis vector whose \(i\)-th coordinate is \(1\) and whose other coordinates are \(0\), every vector \(z \in \mathbb {R}^{\iota }\) satisfies
Let \(X, Y \subseteq \mathbb {R}\) be nonempty convex sets with \(X\) compact, and let \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) be a payoff function that is bounded above and below on \(X \times Y\). Suppose that for every \(y \in Y\) the map \(x \mapsto f(x,y)\) is continuous and convex on \(X\), and that for every \(x \in X\) the map \(y \mapsto f(x,y)\) is concave on \(Y\). Then the minimax identity holds:
Let \(X, Y \subseteq \mathbb {R}\) be nonempty sets with \(X\) compact and both \(X\) and \(Y\) convex, and let \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) be a payoff function that is bounded above and below on \(X \times Y\). Suppose that for every \(y \in Y\) the section \(x \mapsto f(x,y)\) is continuous and convex on \(X\), while for every \(x \in X\) the section \(y \mapsto f(x,y)\) is concave on \(Y\). Then the minimax equality
holds.
Let \(X, Y \subseteq \mathbb {R}\) be nonempty, compact, convex sets, and let \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) be jointly continuous on \(X \times Y\) and normalized so that \(0 \le f(x,y) \le 1\) for all \(x \in X\) and \(y \in Y\). Suppose that \(f\) is bounded above and below on \(X \times Y\), that for every \(y \in Y\) the map \(x \mapsto f(x,y)\) is continuous and convex on \(X\), and that for every \(x \in X\) the map \(y \mapsto f(x,y)\) is concave on \(Y\). Then
Let \(X, Y \subseteq \mathbb {R}\) with \(X\) and \(Y\) nonempty, \(X\) compact and convex, and \(Y\) convex, and let \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) be bounded above and below on \(X \times Y\), with \(f(\cdot , y)\) continuous and convex on \(X\) for every \(y \in Y\) and \(f(x, \cdot )\) concave on \(Y\) for every \(x \in X\). Write \(v = \sup _{y \in Y} \inf _{x \in X} f(x,y)\) for the lower value, and suppose that for every \(\varepsilon \gt 0\) and every finite set \(u\) of column points from \(Y\), the set
is nonempty. Then the minimax equality
holds.
Let \(X, Y \subseteq \mathbb {R}\) and let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\). Suppose that \(X\) and \(Y\) are nonempty, that \(X\) is compact and convex, that \(Y\) is convex, that the map \(x \mapsto f(x,y)\) is continuous and convex on \(X\) for every \(y \in Y\), that the map \(y \mapsto f(x,y)\) is concave on \(Y\) for every \(x \in X\), and that \(f\) is bounded above and below on \(X \times Y\). Suppose moreover that for every nonempty finite subset \(u \subseteq X\),
and that for every \(\varepsilon \gt 0\) there exists a nonempty finite subset \(u \subseteq X\) with
Then
A structure bundling the assumptions of Cesa-Bianchi–Lugosi Theorem 7.1 for \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) on sets \(X, Y \subseteq \mathbb {R}\): both sets are nonempty, \(X\) is compact and convex, \(Y\) is convex, \(f(\cdot ,y)\) is continuous and convex on \(X\) for every \(y \in Y\), \(f(x,\cdot )\) is concave on \(Y\) for every \(x \in X\), and \(f\) is bounded above and below on \(X \times Y\).
For sets \(X, Y \subseteq \mathbb {R}\) and a payoff function \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\), OnlineLearning.ConvexCompactMinimaxStatement is the proposition
It is packaged as a named Prop so that intermediate lemmas can refer to it while the proof is built up incrementally.
Let \(S \subseteq \mathbb {R}\) be a convex set and let \(g : \mathbb {R}\to \mathbb {R}\) be convex on \(S\). Let \(p\) be a mixed strategy over \(n\) actions, that is, a family of weights \(p_1, \dots , p_n \ge 0\) with \(\sum _{i=1}^{n} p_i = 1\), and let \(x_1, \dots , x_n\) be points of \(S\). Then
Let \(X, Y \subseteq \mathbb {R}\) together with a payoff function \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) satisfy the convex-compact minimax hypotheses, and let \(c \in \mathbb {R}\). Suppose that for every finite subset \(u \subseteq Y\) the set \(X \cap \bigcap _{y \in u} \mathrm{sublevel}(X,f,y,c)\) is nonempty, where \(\mathrm{sublevel}(X,f,y,c) = X \cap \{ x \in \mathbb {R}: f(x,y) \le c\} \); that is, there is a common row point \(x \in X\) with \(f(x,y) \le c\) for all \(y\) in that finite set. Then there exists a single \(x \in X\) with \(f(x,y) \le c\) for every \(y \in Y\).
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N\) column actions, whose payoff matrix \(A = (A_{ij})\) satisfies \(0 \le A_{ij} \le 1\) for all \(i,j\). Suppose \(M \gt 1\) and \(N \ge 1\). Then the lower and upper values of \(G\) coincide:
where \(\Delta _M\) and \(\Delta _N\) denote the sets of mixed strategies (probability distributions) over the row and column actions respectively.
Let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff function, and let \(x : [M] \to \mathbb {R}\) and \(y : [N] \to \mathbb {R}\) be finite families of row and column points such that \(0 \le f(x_i, y_j) \le 1\) for all \(i\) and \(j\). Suppose the column action set is nonempty (\(N \ge 1\)) and there are at least two row actions (\(M \ge 2\)). Then the sampled game — the \(M \times N\) matrix game with entries \(f(x_i, y_j)\) — has equal lower and upper values,
where \(\Delta _M\) and \(\Delta _N\) denote the sets of mixed strategies over the row and column actions.
Let \(X, Y \subseteq \mathbb {R}\) and \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) satisfy the convex-compact minimax hypotheses, and write \(v = \sup _{y \in Y} \inf _{x \in X} f(x,y)\) for the lower value of the payoff. Then for every \(\varepsilon \gt 0\) and every finite set \(u\) of columns chosen from \(Y\), there is a row \(x \in X\) with \(f(x,y) \le v + \varepsilon \) simultaneously for all \(y \in u\); equivalently, the intersection over \(y \in u\) of the sublevel sets \(\{ x \in X : f(x,y) \le v + \varepsilon \} \) is nonempty.
Let \(X, Y \subseteq \mathbb {R}\) and \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) satisfy the convex-compact minimax hypotheses, and suppose \(f\) is normalized so that \(0 \le f(x,y) \le 1\) for all \(x \in X\) and \(y \in Y\). Fix an integer \(M \ge 2\), a row sample \(x : \mathrm{Fin}\, M \to X\), a real \(\varepsilon \gt 0\), and a finite column sample \(v \subseteq Y\). Write
for the indexed sampled lower value. Then the set
is nonempty; that is, there is a row \(\bar{x} \in X\) with \(f(\bar{x}, y) \le K + \varepsilon \) for every \(y \in v\). This is the finite \(\varepsilon \)-net step in the proof of Cesa-Bianchi–Lugosi Theorem 7.1.
For a row sample indexed by \(\mathrm{Fin}\, M\), the finite indexed row sample lower value is
where \(x : \mathrm{Fin}\, M \to X\) is the sample. This is a technical variant of \(\texttt{OnlineLearning.finiteRowSampleLowerValue}\) used in intermediate results.
For a finite zero-sum game \(G\) with \(M\) row actions and \(N\) column actions, the lower value is
where the row player maximises over mixed strategies and the column player minimises over pure responses.
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N \ge 1\) column actions, whose payoff matrix \((A_{ij})\) satisfies \(0 \le A_{ij} \le 1\) for all \(i,j\). For a mixed row strategy \(p = (p_i)\) over the \(M\) rows, let \(L(p) = \inf _{j} \sum _i p_i A_{ij}\) be the least expected payoff against a pure column, taken over all \(N\) columns. Then the set \(\{ \, L(p) : p \text{ is a mixed strategy over the rows} \, \} \) is bounded above.
Let \(G\) be a finite two-player zero-sum game with \(M \geq 1\) row actions and \(N \geq 1\) column actions, given by a payoff matrix \(A_{ij} \in [0,1]\) that the row player seeks to maximise and the column player to minimise. Then the lower value never exceeds the upper value:
where \(\Delta _M\) and \(\Delta _N\) denote the sets of mixed strategies (probability distributions) over the row and column actions, respectively.
Given a payoff \(f : \mathbb {R} \to \mathbb {R} \to \mathbb {R}\) and a finite subset \(u \subseteq X\), the finite row sample lower value is
i.e. the lower value obtained when the row player is restricted to the finite sample \(u\) while the column player still ranges over all of \(Y\).
For sets \(X, Y \subseteq \mathbb {R}\), a payoff \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\), and a finite sample \(u\) of columns from \(Y\), the finite upper image is the set
of vectors that strictly dominate the payoff vector of some row on the sample \(u\). This set is written \(\texttt{OnlineLearning.finiteUpperImage}\) in Lean.
Let \(X, Y \subseteq \mathbb {R}\), let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff, and let \(u\) be a finite sample of columns from \(Y\). If \(X\) is convex and, for every \(y \in Y\), the map \(x \mapsto f(x,y)\) is convex on \(X\), then the finite upper image \(U(X,f,u)\) is a convex subset of \(u \to \mathbb {R}\).
Let \(X, Y \subseteq \mathbb {R}\), let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff function, and let \(u\) be a finite sample of columns from \(Y\). Then the finite upper image
is an open subset of the space \(u \to \mathbb {R}\) of vectors indexed by \(u\).
Let \(X, Y \subseteq \mathbb {R}\), let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff function, and let \(u\) be a finite sample of columns from \(Y\). If \(X\) is nonempty, then the finite upper image \(U(X,f,u)\) is nonempty; that is, there exists a vector \(z : u \to \mathbb {R}\) for which some \(x \in X\) satisfies \(f(x,y) \lt z_y\) for every \(y \in u\).
Let \(X, Y \subseteq \mathbb {R}\), let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff, and let \(u\) be a finite sample of columns from \(Y\). If a vector \(z : u \to \mathbb {R}\) belongs to the finite upper image \(U(X,f,u)\), and \(z' : u \to \mathbb {R}\) satisfies \(z_y \le z'_y\) for every \(y \in u\), then \(z'\) also belongs to \(U(X,f,u)\).
Let \(G\) be a finite two-player zero-sum game with \(M \ge 1\) row actions and \(N\) column actions, whose payoff matrix \(A\) satisfies \(A_{ij} \in [0,1]\) for all \(i,j\). For a mixed column strategy \(q\) (a probability distribution on the \(N\) columns), let
be the best expected payoff that a single pure row can achieve against \(q\). Then the set \(\{ \, v(q) : q \text{ a mixed column strategy}\, \} \) is bounded below.
Let \(X \subseteq \mathbb {R}\) be a closed set, let \(f\colon \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) be a payoff function, and fix a column point \(y \in \mathbb {R}\) and a threshold \(c \in \mathbb {R}\). If the map \(x \mapsto f(x,y)\) is continuous on \(X\), then the minimax sublevel set \(\mathrm{sublevel}(X,f,y,c) = X \cap \{ x \in \mathbb {R}: f(x,y) \le c\} \) is closed.
Let \(X \subseteq \mathbb {R}\) be compact, fix a payoff function \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\), a column point \(y \in \mathbb {R}\), and a threshold \(c \in \mathbb {R}\), and suppose the map \(x \mapsto f(x,y)\) is continuous on \(X\). Then the minimax sublevel set \(\mathrm{sublevel}(X,f,y,c) = X \cap \{ x \in \mathbb {R}: f(x,y) \le c\} \) is compact.
Let \(X \subseteq \mathbb {R}\) be a set, let \(f \colon \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) be a payoff function, and let \(y, c \in \mathbb {R}\). Then the minimax sublevel set \(\mathrm{sublevel}(X,f,y,c) = X \cap \{ x \in \mathbb {R}: f(x,y) \leq c\} \) is contained in \(X\).
Let \(S \subseteq \mathbb {R}\) be a convex set, let \(p\) be a mixed strategy over \(n\) actions with weights \(p_1, \dots , p_n \ge 0\) satisfying \(\sum _{i=1}^{n} p_i = 1\), and let \(x_1, \dots , x_n \in S\). Then the weighted average \(\sum _{i=1}^{n} p_i\, x_i\) again lies in \(S\).
Let \(G\) be a finite two-player zero-sum game whose payoff matrix \(A\) has every entry in \([0,1]\), let \(p\) be a mixed row strategy with weights \(p_i\), and let \(j\) be a fixed pure column action. Then the row player’s expected payoff against column \(j\) satisfies
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N\) column actions, let \(p\) be a mixed row strategy over the \(M\) rows, and let \(j\) be a pure column action. Then the expected payoff of \(p\) against column \(j\) is non-negative:
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N\) column actions, whose payoff matrix \(A\) satisfies \(0 \le A_{ij} \le 1\) for all entries, and let \(q\) be a mixed column strategy, that is, a probability distribution \(q_1, \dots , q_N\) over the columns. Then for every pure row action \(i\), the expected payoff of row \(i\) against \(q\) satisfies
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N\) column actions, whose payoff matrix \(A\) has entries satisfying \(0 \le A_{ij} \le 1\). Fix a pure row action \(i\) and a mixed column strategy \(q\), so its weights \(q_j\) are non-negative and sum to \(1\). Then the expected payoff of row \(i\) against \(q\) is non-negative:
Let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) take values in \([0,1]\), and let \(x : [M] \to \mathbb {R}\) and \(y : [N] \to \mathbb {R}\) be finite families of row and column points, giving rise to the sampled game whose \((i,j)\) entry is \(f(x_i, y_j)\). Fix a column index \(j\), and suppose the section \(x' \mapsto f(x', y_j)\) is convex on a set \(X \subseteq \mathbb {R}\) with \(x_i \in X\) for every \(i\). Then for any mixed row strategy \(p\) over the \(M\) rows,
where the right-hand side is the expected payoff of the sampled game against column \(j\) under \(p\).
Let \(f:\mathbb {R}\times \mathbb {R}\to \mathbb {R}\) be a payoff function taking values in \([0,1]\), and let \(x:[M]\to \mathbb {R}\) and \(y:[N]\to \mathbb {R}\) be finite families of row and column points. Fix a row index \(i\) and a mixed column strategy \(q\), whose weights \(q_j\ge 0\) sum to \(1\). Suppose \(Y\subseteq \mathbb {R}\) contains every \(y_j\) and the section \(y'\mapsto f(x_i,y')\) is concave on \(Y\). Then the expected payoff of pure row \(i\) against \(q\) in the sampled game is at most the value of \(f\) at the \(q\)-averaged column point:
Given a payoff function \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) with values in \([0,1]\), finite families of row points \(x : [M] \to \mathbb {R}\) and column points \(y : [N] \to \mathbb {R}\), the sampled game is the \(M\times N\) matrix game whose \((i,j)\) entry is \(f(x_i, y_j)\).
Let \(X, Y \subseteq \mathbb {R}\), let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff, let \(u\) be a finite sample of columns from \(Y\), and let \(c : u \to \mathbb {R}\) be a fixed vector. Suppose \(X\) is nonempty and \(L : (u \to \mathbb {R}) \to \mathbb {R}\) is a continuous linear functional such that \(L(z) \lt L(c)\) for every \(z\) in the finite upper image \(U(X,f,u)\). Then for each coordinate \(y \in u\) one has \(L(e_y) \le 0\), where \(e_y \in (u \to \mathbb {R})\) is the standard basis vector taking the value \(1\) at \(y\) and \(0\) elsewhere.
Let \(X, Y \subseteq \mathbb {R}\), let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff, and let \(u\) be a finite sample of columns drawn from \(Y\), with \(X\) nonempty. Fix a vector \(c : u \to \mathbb {R}\) and a continuous linear functional \(L : (u \to \mathbb {R}) \to \mathbb {R}\). If \(L(z) \lt L(c)\) for every \(z\) in the finite upper image \(U(X,f,u)\), then \(L\) is not the zero functional.
Let \(X, Y \subseteq \mathbb {R}\), let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff, let \(u\) be a finite sample of columns from \(Y\), and let \(c \in \mathbb {R}\). Suppose \(L : \mathbb {R}^u \to \mathbb {R}\) is a continuous linear functional separating the constant vector \((c,\dots ,c)\) from the finite upper image, i.e. \(L(z) \lt L(c,\dots ,c)\) for every \(z \in U(X,f,u)\), and that the total weight \(W = \sum _{y \in u} \bigl(-L(e_y)\bigr)\) is strictly positive, where \(e_y\) is the standard basis vector at \(y\). Then for every row \(x \in X\),
that is, \(c\) is at most the payoff of \(x\) averaged under the normalized weights \(-L(e_y)/W\).
Let \(X, Y \subseteq \mathbb {R}\), let \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) be a payoff function, and let \(u\) be a finite sample of columns drawn from \(Y\). Suppose \(X\) is nonempty, fix a vector \(c \in \mathbb {R}^u\), and let \(L : \mathbb {R}^u \to \mathbb {R}\) be a continuous linear functional that strictly separates \(c\) from the finite upper image \(U(X,f,u)\), in the sense that \(L(z) \lt L(c)\) for every \(z \in U(X,f,u)\). Writing \(e_y \in \mathbb {R}^u\) for the standard basis vector supported at coordinate \(y\), one then has
Let \(X, Y \subseteq \mathbb {R}\) and \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) satisfy the convex-compact minimax hypotheses, and suppose in addition that \(Y\) is compact and that \(f\) is jointly continuous on \(X \times Y\). Then for every \(\varepsilon \gt 0\) there is a nonempty finite subset \(u \subseteq X\) such that
that is, restricting the row player to a suitable finite sample raises the lower value by at most \(\varepsilon \).
Let \(X, Y \subseteq \mathbb {R}\) be nonempty, with \(X\) compact and convex and \(Y\) convex, and let \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) be bounded above and below on \(X \times Y\), with \(f(\cdot , y)\) continuous and convex on \(X\) for every \(y \in Y\) and \(f(x, \cdot )\) concave on \(Y\) for every \(x \in X\). Suppose in addition that \(f\) is uniformly equicontinuous in its first argument, uniformly over the columns: for every \(\varepsilon \gt 0\) there is \(\rho \gt 0\) with \(|f(x,y) - f(x',y)| \lt \varepsilon \) whenever \(x, x' \in X\), \(y \in Y\), and \(\operatorname{dist}(x, x') \lt \rho \). Then for every \(\varepsilon \gt 0\) there exists a nonempty finite sample \(u \subseteq X\) such that
that is, restricting the row player to a suitable finite subset of \(X\) raises the lower value of the game by at most \(\varepsilon \).
Let \(X, Y \subseteq \mathbb {R}\) satisfy the convex-compact minimax hypotheses for a payoff \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\), and suppose in addition that \(f\) is normalized, \(0 \le f(x,y) \le 1\) for all \(x \in X\) and \(y \in Y\). Fix an integer \(M \ge 2\) and a finite row sample \(x : \mathrm{Fin}\, M \to X\), writing \(x_i\) for its \(i\)-th point. Then
Let \(X, Y \subseteq \mathbb {R}\) and \(f : \mathbb {R}\to \mathbb {R}\to \mathbb {R}\) satisfy the convex-compact minimax hypotheses, and suppose the payoff is normalized so that \(0 \le f(x,y) \le 1\) for all \(x \in X\) and \(y \in Y\). Then for every nonempty finite subset \(u \subseteq X\),
that is, the game’s upper value is bounded above by the lower value of the game in which the row player is restricted to the finite sample \(u\) while the column player still ranges over all of \(Y\).
Let \(X, Y \subseteq \mathbb {R}\) and \(f : \mathbb {R}\times \mathbb {R}\to \mathbb {R}\) satisfy the convex-compact minimax hypotheses, and suppose in addition that \(Y\) is compact and that \((x,y) \mapsto f(x,y)\) is continuous on \(X \times Y\). Then \(f\) is uniformly equicontinuous in its first argument, uniformly over all columns: for every \(\varepsilon \gt 0\) there exists \(\rho \gt 0\) such that for all \(x, x' \in X\) and all \(y \in Y\),
\(\texttt{OnlineLearning.Theorem71CompactApproximation}\) is the proposition asserting that for every \(\varepsilon \gt 0\) there exists a nonempty finite sample \(u \subseteq X\) such that
This encodes the “let the net size go to zero” compactness step of the textbook proof.
\(\texttt{OnlineLearning.Theorem71FiniteNoRegretBound}\) is the proposition asserting that for every nonempty finite row sample \(u \subseteq X\),
This is the finite-game no-regret ingredient in the Theorem 7.1 proof route.
\(\texttt{OnlineLearning.Theorem71UniformEquicontinuity}\) holds for \(X\), \(Y\), \(f\) when \(f\) is uniformly equicontinuous in its first argument uniformly over all columns: for every \(\varepsilon \gt 0\) there exists \(\rho \gt 0\) such that for all \(x, x' \in X\) and \(y \in Y\),
Let \(X, Y \subseteq \mathbb {R}\) be nonempty convex sets with \(X\) also compact, and let \(f : X \times Y \to \mathbb {R}\) be a payoff that is bounded above and below on \(X \times Y\), such that for every \(y \in Y\) the map \(x \mapsto f(x,y)\) is continuous and convex on \(X\), and for every \(x \in X\) the map \(y \mapsto f(x,y)\) is concave on \(Y\). Then
For real numbers \(a, T, c \gt 0\), at the learning rate \(\eta = \sqrt{c a / T}\) the two terms \(a/\eta \) and \(\eta T / c\) sum to \(\sqrt{4 a T / c}\):
Fix a prime \(p\) and work in the symplectic space \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) equipped with the form \(\omega \big((x,z),(x',z')\big) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\), and for a subset \(M \subseteq \{ 0,\dots ,n-1\} \) let \(V_M\) be the submodule of vectors whose support lies in \(M\) and let \(r_M \colon V \to V\) be the map that zeroes every coordinate outside \(M\) and fixes those in \(M\). Then for every \(\mathbb {F}_p\)-subspace \(S \le V\),
where \(r_M(S)\) denotes the image of \(S\) under \(r_M\) and \((\cdot )^{\perp _\omega }\) the orthogonal complement with respect to \(\omega \).
Let \(p\) be a prime and let \(\omega \) be the symplectic form on \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\). For any subset \(M\) of \(\{ 0, 1, \dots , n-1\} \) and any submodule \(S \le V\), the vectors of the support submodule \(V_M\) that are \(\omega \)-orthogonal to every element of \(S\) are exactly the images, under the inclusion \(V_M \hookrightarrow V\), of the vectors of \(V_M\) that are orthogonal to \(r_M(S)\) with respect to the restriction of \(\omega \) to \(V_M\). In symbols, writing \(\iota \colon V_M \hookrightarrow V\) for the inclusion and \(\perp _M\) for orthogonality with respect to the restricted form,
where \(r_M(S)\) is the image of \(S\) under the restriction map \(r_M \colon V \to V_M\).
The three-element type \(\{ X, Y, Z\} \) of labels for the non-identity single-qubit Pauli operators. It serves as the index set in counting arguments that enumerate, for a fixed support, the \(3^{|S|}\) assignments of a non-identity Pauli to each supported coordinate.
Let \(G\) be a finite two-player zero-sum game with \(M \ge 1\) row actions and \(N\) column actions, given by a payoff matrix \(A\) with \(0 \le A_{ij} \le 1\), and fix a pure column action \(j\). Let \(T \gt 0\) and let \(p_0, \dots , p_{T-1}\) be probability distributions over the \(M\) row actions, each presented as non-negative weights \(p_t(i)\) summing to \(1\), and let \(\bar{p}\) be their average strategy, the mixed row strategy with weights \(\bar{p}_i = \frac{1}{T}\sum _{t=0}^{T-1} p_t(i)\). Then the expected payoff of \(\bar{p}\) against column \(j\) satisfies
Fix \(N \ge 1\) experts and \(T\) rounds, a loss sequence \(\ell \), and a learning rate \(\eta \gt 0\). Then the final potential is at least the unnormalized Hedge weight of the best expert in hindsight:
where \(L^* = \min _{i} L_T(i)\) is the smallest cumulative loss over the full horizon \(T\).
Fix a learning rate \(\eta \in \mathbb {R}\), a loss sequence \(\ell \) for \(N \ge 1\) experts over \(T\) rounds, and a time \(t\). Then the potential at time \(t\) is strictly positive:
where \(L_t(i) = \sum _{s \lt t} \ell _s(i)\) is the cumulative loss of expert \(i\) through the first \(t\) rounds.
Fix \(N \ge 1\) experts, a horizon of \(T\) rounds, a learning rate \(\eta \gt 0\), and a valid loss sequence \(\ell \). Then for every round \(t\) with \(t \lt T\), the consecutive potentials satisfy
where \(\widehat{\ell }_t = \sum _{i=1}^{N} p_t(i)\, \ell _t(i)\) is the expected loss of the learner under the Hedge distribution at round \(t\).
Fix \(N \ge 1\) experts, a horizon \(T\), and a loss sequence \(\ell \) over these experts and rounds, and let \(\eta \in \mathbb {R}\) be a learning rate. At time \(0\) every expert has zero cumulative loss, so each unnormalized Hedge weight equals \(1\), and the potential is
Given a game \(G\), a prefix of \(t\) column actions, and a row \(i\), the prefix cumulative loss of row \(i\) is
This mirrors cumLoss from the abstract Hedge framework, expressed directly in terms of game payoffs.
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N\) column actions, given by a payoff matrix \(A\) with \(0 \le A_{ij} \le 1\), and let \(a_0, \dots , a_{T-1}\) be a sequence of column actions with each \(a_s \in \mathrm{Fin}\, N\). Fix a row action \(i \in \mathrm{Fin}\, M\) and a round \(t \in \mathrm{Fin}\, T\). Then the prefix cumulative loss of row \(i\) along the first \(t\) column actions coincides with the cumulative loss of expert \(i\) through the first \(t\) rounds of the loss sequence induced by \(G\) and \((a_s)_s\); that is,
Let \(G\) be a finite two-player zero-sum game with \(M \ge 1\) row actions and \(N\) column actions, fix a learning rate \(\eta \in \mathbb {R}\), and let \(a_0, \dots , a_{T-1}\) be a sequence of column actions. For any round \(t \lt T\) and any row \(i\), the weight assigned to \(i\) by the prefix Hedge mixed strategy built from the length-\(t\) prefix \(a_0, \dots , a_{t-1}\) equals the Hedge distribution value \(p_t(i)\) obtained by running abstract Hedge, at the same learning rate \(\eta \), on the loss sequence \(\ell \) induced by \(G\) and \(a\), namely \(\ell _s(i) = 1 - A_{i, a_s}\). In symbols, both sides equal \(w_t(i)/\Phi _t\), where \(w_t(i) = \exp (-\eta \sum _{s\lt t}(1 - A_{i,a_s}))\) and \(\Phi _t = \sum _{k} w_t(k)\).
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N\) column actions, with payoff matrix \(A\), let \(\eta \in \mathbb {R}\) be a learning rate, and let \(a_0, \dots , a_{T-1}\) be a sequence of column actions, so that \(a : \mathrm{Fin}\, T \to \mathrm{Fin}\, N\). Fix a round \(t \in \mathrm{Fin}\, T\) and a row \(i \in \mathrm{Fin}\, M\). Then the prefix Hedge weight of row \(i\) at learning rate \(\eta \), computed from the length-\(t\) prefix \(a_0, \dots , a_{t-1}\) of column actions, equals the Hedge weight of expert \(i\) at time \(t\) and learning rate \(\eta \) for the loss sequence \(\ell _s(j) = 1 - A_{j, a_s}\) induced by \(G\) along \(a\); that is, both equal
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N\) column actions, let \(\eta \in \mathbb {R}\) be a learning rate, and fix any prefix \(a_0, \dots , a_{t-1}\) of \(t\) column actions together with a row \(i\). Then the prefix Hedge weight \(w_t(i) = \exp \! \bigl(-\eta \cdot L_t(i)\bigr)\) is strictly positive, where \(L_t(i) = \sum _{s=0}^{t-1}\bigl(1 - A_{i,a_s}\bigr)\) is the prefix cumulative loss of row \(i\).
Let \(G\) be a finite two-player zero-sum game with payoff matrix \(A\) having \(M\) rows and \(N\) columns, fix a learning rate \(\eta \in \mathbb {R}\), and let \(a_0, \dots , a_{T-1}\) be a sequence of \(T\) column actions. For any round \(t \lt T\), the prefix potential of \(G\) at learning rate \(\eta \) formed from the first \(t\) column actions \(a_0, \dots , a_{t-1}\) equals the Hedge potential at time \(t\) of the loss sequence \(\ell _s(i) = 1 - A_{i, a_s}\) induced on the row player by \(G\) and the full sequence of column actions; that is,
where \(L_t(i) = \sum _{s \lt t} \ell _s(i)\) is the cumulative loss of row \(i\).
Let \(G\) be a finite two-player zero-sum game with at least one row action, fix a learning rate \(\eta \in \mathbb {R}\), and let \(a_0, \dots , a_{t-1}\) be any prefix of column actions. Then the prefix potential is strictly positive,
where \(L_t(i)\) is the prefix cumulative loss of row \(i\).
For a boolean-vector-valued random variable \(X : \Omega \to (\mathrm{Fin}\, m \to \mathrm{Bool})\) and an index \(i : \mathrm{Fin}\, m\), the strict prefix \(\mathrm{boolVectorStrictPrefix}\, X\, i\, \omega \) is the restriction of \(X(\omega )\) to coordinates \(0, 1, \ldots , i-1\).
Let \((Y_i)_{i}\) be a finite family of measurable real random variables on a probability space \((\Omega , \mu )\), indexed by a finite set of cardinality \(n\), and suppose the \(Y_i\) are mutually independent. Fix real numbers \(\sigma ^2\), \(c \gt 0\), and \(t_{\max }\), and suppose that for each \(i\) the centered square \(Y_i^2 - \sigma ^2\) satisfies the Bernstein MGF condition with parameters \(c\) and \(t_{\max }\). Then for every \(s\) with \(0 \le s \le 2 n c\, t_{\max }\),
Let \(\mu \) be a finite measure on a measurable space, and let \(F \subseteq A\) be measurable sets. Then conditioning \(\mu \) on \(A\) and afterwards on \(F\) gives the same measure as conditioning \(\mu \) on \(F\) directly:
Let \((\Omega ,\mu )\) be a probability space, and let \(X,X'\colon \Omega \to S\) and \(Y,Y'\colon \Omega \to T\) be measurable random variables valued in countable measurable spaces, each taking only finitely many values. If \(X = X'\) and \(Y = Y'\) hold \(\mu \)-almost everywhere, then the conditional Shannon entropies agree:
Let \(Y\) and \(Z\) be random variables on a probability space \((\Omega , \mu )\) (allowing also the zero measure), each taking countably many values with finite range, and let \(X : \Omega \to (\{ 0,1\} ^m)\) be a random boolean vector; assume \(X\), \(Y\), and \(Z\) are measurable. Writing \(X_i(\omega ) = X(\omega )(i)\) for the \(i\)-th coordinate and \(X_{\lt i}(\omega )\) for the strict prefix consisting of coordinates \(0, 1, \ldots , i-1\), the conditional mutual information of \(X\) and \(Y\) given \(Z\) decomposes as
Let \(X\), \(Y\), and \(W\) be random variables on a probability space \((\Omega ,\mu )\), taking values in countable measurable spaces and each attaining only finitely many values, and let \(f\) be a measurable map defined on the range of \(W\). Suppose the conditional mutual information \(I[X : W \mid f\circ W;\mu ]\) vanishes, so that \(X\) and \(W\) are conditionally independent given the coarsened variable \(f\circ W\). Then
Let \(X\), \(Y\), and \(Z\) be random variables on a probability space \((\Omega , \mu )\), taking values in countable measurable spaces \(S\), \(T\), and \(U\) respectively, each measurable and of finite range. For any function \(f \colon U \times S \to V\) into a countable measurable space \(V\), form the postprocessed variable \(W(\omega ) = f\bigl(Z(\omega ), X(\omega )\bigr)\), obtained by applying to \(X\) a map that may itself depend on the conditioning value \(Z\). Then the conditional mutual information satisfies
Let \(\mu \) be a probability measure on a sample space \(\Omega \), and let \(X\), \(Y\), \(Z\) be measurable random variables on \(\Omega \), taking values in countable measurable spaces \(S\), \(T\), \(U\) respectively and each having finite range. Let \(f\colon T\to V\) and \(g\colon U\to W\) be measurable injections into countable measurable spaces. Then relabelling the second and conditioning variables by \(f\) and \(g\) leaves the conditional mutual information unchanged:
Let \(X\colon \Omega \to S\), \(Y\colon \Omega \to T\), and \(Z\colon \Omega \to U\) be measurable random variables on a probability space \((\Omega ,\mu )\), each taking only countably many values and having finite range, where \(S\), \(T\), and \(U\) are countable measurable spaces. Let \(f\colon U\times T\to V\) be a measurable map into a countable measurable space \(V\), and define \(W(\omega )=f\big(Z(\omega ),Y(\omega )\big)\). Then the conditional mutual information satisfies
That is, replacing \(Y\) by any function of \(Y\) that is also permitted to depend on the conditioning variable \(Z\) cannot increase the conditional mutual information with \(X\).
Let \(X\), \(Y\), and \(Z\) be discrete random variables on a probability space \((\Omega , \mu )\), each taking values in a countable set and attaining only finitely many values, and let \(f\) be an arbitrary function defined on the range of \(Z\). Then the conditional mutual information of \(X\) and \(Y\) given the augmented variable \((Z, f(Z))\) equals the conditional mutual information of \(X\) and \(Y\) given \(Z\) alone:
Let \(X\), \(Y\), \(Z\) be random variables on a probability space \((\Omega ,\mu )\) (where \(\mu \) is allowed to be the zero measure), taking values in countable measurable spaces \(S\), \(T\), \(U\) respectively, and assume each of \(X\), \(Y\), \(Z\) is measurable with finite range. Let \(V\) be a further countable measurable space and let \(f\colon U\times S\to V\) be measurable, and write \(f(Z,X)\) for the random variable \(\omega \mapsto f\bigl(Z(\omega ),X(\omega )\bigr)\). Then conditioning additionally on \(f(Z,X)\) cannot increase the conditional mutual information of \(X\) and \(Y\) given \(Z\):
Let \(X\), \(Y\), \(Z\) be random variables on a probability space \((\Omega , \mu )\), taking values in countable measurable spaces, each measurable and with finite range. Let \(f\) be a measurable function of two arguments, and write \(f(Z, Y)\) for the random variable \(\omega \mapsto f\bigl(Z(\omega ), Y(\omega )\bigr)\). Then conditioning on the pair \(\bigl(Z, f(Z, Y)\bigr)\) — that is, refining the conditioning variable \(Z\) by a deterministic function of \(Z\) and \(Y\) — cannot increase the conditional mutual information between \(X\) and \(Y\):
Let \((\Omega ,\mu )\) be a probability space, and let \(X,X'\), \(Y,Y'\), and \(Z,Z'\) be measurable random variables on \(\Omega \) taking values in countable (discrete) measurable spaces, each attaining only finitely many values. If \(X = X'\), \(Y = Y'\), and \(Z = Z'\) hold \(\mu \)-almost everywhere, then the conditional mutual informations agree:
Let \(\mu \) be a probability measure on a finite sample space \(\Omega \) (with every subset measurable), and let \(X, X'\), \(Y, Y'\), and \(Z, Z'\) be random variables on \(\Omega \) taking values in the finite sets \(S\), \(T\), and \(U\) respectively. If \(X = X'\), \(Y = Y'\), and \(Z = Z'\) each hold \(\mu \)-almost everywhere, then the conditional mutual informations coincide:
Let \(\mu \) be a probability measure on a sample space \(\Omega \), and let \(X, X' \colon \Omega \to S\), \(Y, Y' \colon \Omega \to T\), and \(Z \colon \Omega \to U\) be measurable maps into countable spaces, each taking only finitely many values \(\mu \)-almost surely. If \(X = X'\) and \(Y = Y'\) hold \(\mu \)-almost everywhere, then the conditional mutual informations agree:
Let \(\mu \) be a probability measure on a sample space \(\Omega \) (or the zero measure), and let \(X\), \(Y\), \(Z\) be measurable random variables on \(\Omega \) taking finitely many values in the countable sets \(S\), \(T\), \(U\) respectively. If there is a constant \(c \in S\) with \(X = c\) holding \(\mu \)-almost everywhere, then the conditional mutual information of \(X\) and \(Y\) given \(Z\) vanishes:
Let \(X\), \(Y\), and \(Z\) be discrete random variables on a probability space \((\Omega , \mu )\) — where \(\mu \) is either the zero measure or a probability measure — taking values in countable measurable spaces \(S\), \(T\), and \(U\) respectively, with each of \(X\), \(Y\), \(Z\) almost surely taking only finitely many values, and with \(X\) and \(Y\) measurable. If there is a constant \(c \in T\) such that \(Y = c\) \(\mu \)-almost everywhere, then the conditional mutual information satisfies \(I[X : Y \mid Z; \mu ] = 0\).
Let \(X\), \(Y\), and \(Z\) be random variables on a probability space \((\Omega , \mu )\), taking values in countable measurable spaces, and assume each takes only finitely many values (with positive probability). Then the conditional mutual information of \(X\) and \(Y\) given \(Z\) is at most the conditional entropy of \(X\) given \(Z\):
Here \(\mu \) is allowed to be either a probability measure or the zero measure.
Let \(X\), \(Y\), and \(Z\) be measurable random variables on a common probability space \((\Omega , \mu )\), each taking values in a countable measurable space and each having finite range, and suppose \(\mu \) is either the zero measure or a probability measure. Then the conditional mutual information of \(X\) and \(Y\) given \(Z\) is at most the conditional entropy of \(Y\) given \(Z\):
Let \(X\), \(Y\), and \(Z\) be measurable random variables on a probability space \((\Omega , \mu )\), taking values in countable spaces and each attaining only finitely many values. Then the conditional mutual information of \(X\) and \(Y\) given \(Z\) is bounded above by the Shannon entropy of \(X\):
(The bound also holds trivially when \(\mu \) is the zero measure.)
Let \(X\), \(Y\), and \(Z\) be random variables on a probability space \((\Omega , \mu )\), taking values in countable measurable spaces and each attaining only finitely many values. Then the mutual information between \(X\) and \(Y\) conditioned on \(Z\) is at most the Shannon entropy of \(Y\):
Let \(X\), \(Y\), and \(W\) be random variables on a probability space \((\Omega ,\mu )\) (with \(\mu \) a probability measure, or else the zero measure), taking values in countable measurable spaces and each assuming only finitely many values. Write \(I[X : Y;\mu ]\) for the mutual information of \(X\) and \(Y\), and \(I[X : Y \mid W;\mu ]\) for their mutual information conditioned on \(W\). If conditioning on \(Y\) does not increase the mutual information between \(X\) and \(W\), that is,
then conditioning on \(W\) does not increase the mutual information between \(X\) and \(Y\):
Let \(X\), \(Y\), \(W\), and \(Z\) be measurable functions on a probability space \((\Omega , \mu )\), taking values in countable measurable spaces and each having finite range. Then the conditional mutual information satisfies
where \((Y, W)\) denotes the joint random variable \(\omega \mapsto (Y(\omega ), W(\omega ))\).
Let \((\Omega ,\mu )\) be a finite measure space, and let \(X\), \(Y\), \(K\), and \(Z\) be measurable functions on \(\Omega \) taking values in countable spaces, with \(X\) and \(Y\) each having finite range, \(K\) taking values in a finite set \(A\), and \(Z\) taking values in a finite space. Then the mutual information of \(X\) and \(Y\) conditioned on the pair \((K,Z)\) decomposes as the \(\mu \)-weighted average, over the values of \(K\), of the mutual information of \(X\) and \(Y\) conditioned on \(Z\) alone:
where \(\mu (K^{-1}\{ k\} )\) denotes the (real-valued) measure of the event \(\{ K = k\} \) and \(\mu [\, \cdot \mid K = k\, ]\) denotes \(\mu \) conditioned on that event.
Let \(X\), \(W\), \(Y\), and \(Z\) be random variables on a measure space \((\Omega ,\mu )\), each valued in a countable measurable space and each attaining only finitely many values, and suppose \(\mu \) is either the zero measure or a probability measure. Writing \(I[U : V \mid T;\mu ]\) for the mutual information of \(U\) and \(V\) conditioned on \(T\), the pair \((X,W)\) obeys the chain rule
Let \(X\), \(Y\), \(Z\), and \(W\) be measurable functions on a probability space \((\Omega , \mu )\), where \(\mu \) is either the zero measure or a probability measure, each taking values in a countable measurable space and attaining only finitely many values. Write \(I(X ; Y \mid Z)\) for the conditional mutual information between \(X\) and \(Y\) given \(Z\). Then the conditional mutual information between \(X\) and the pair \((Y, W)\) given \(Z\) satisfies
Let \(S\) be a finite nonempty set, let \((\Omega ,\mu )\) be a measure space, and let \(X\colon \Omega \to S\) be a random variable. If \(N\) is a natural number with \(\left\lvert S\right\rvert \le N\), then the Shannon entropy of \(X\) satisfies
Let \(S\) be a finite nonempty set and let \(X\) be a random variable on a probability space \(\Omega \) (equipped with a measure \(\mu \)) taking values in \(S\). If the number of elements of \(S\) is at most \(2^c\) for some natural number \(c\), then the Shannon entropy of \(X\) satisfies
where the logarithm is natural (so the entropy is measured in nats).
\(\texttt{ProbabilityTheory.HasBernsteinMGF}\ X\ \mu \ c\ t_{\max }\) is a predicate on a random variable \(X : \Omega \to \mathbb {R}\), a measure \(\mu \), and real parameters \(c, t_{\max } \ge 0\). It asserts that for every \(t\) with \(|t| \le t_{\max }\), the exponential moment \(e^{tX}\) is \(\mu \)-integrable, and the moment generating function satisfies \(\mathbf{E}_\mu [e^{tX}] \le e^{c t^2}\). The pair \((c, t_{\max })\) parametrizes a “two-parameter sub-exponential” family: \(c\) controls the quadratic growth near zero and \(t_{\max }\) bounds the radius of validity.
Let \(X : \Omega \to \mathbb {R}\) be a random variable on a probability space \((\Omega , \mu )\), and let \(c, t_{\max } \ge 0\) be real parameters such that \(X\) satisfies the Bernstein moment generating function condition with parameters \(c\) and \(t_{\max }\): for every \(t\) with \(\left\lvert t\right\rvert \le t_{\max }\) the exponential moment \(e^{tX}\) is \(\mu \)-integrable and \(\mathbb {E}_\mu [e^{tX}] \le e^{c t^2}\). If \(c \gt 0\), then for every \(s\) with \(0 \le s \le 2 c\, t_{\max }\),
Let \(X : \Omega \to \mathbb {R}\) be a random variable on a finite measure space \((\Omega , \mu )\), and suppose \(X\) satisfies the Bernstein moment-generating-function condition with parameters \(c\) and \(t_{\max }\): for every \(t\) with \(\left\lvert t\right\rvert \le t_{\max }\) the exponential moment \(e^{tX}\) is \(\mu \)-integrable and \(\mathbb {E}_\mu [e^{tX}] \le e^{c t^2}\). If \(c \gt 0\), then for every \(s\) with \(0 \le s \le 2c\, t_{\max }\) the upper tail obeys
Let \(X : \Omega \to \mathbb {R}\) be a random variable on a probability space carrying a finite measure \(\mu \), and suppose \(X\) satisfies the Bernstein moment generating function condition with parameters \(c\) and \(t_{\max }\): for every \(t\) with \(|t| \le t_{\max }\), the exponential moment \(e^{tX}\) is \(\mu \)-integrable and \(\mathbb {E}_\mu [e^{tX}] \le e^{c t^2}\). If \(c \gt 0\), then for every \(s\) with \(0 \le s \le 2 c\, t_{\max }\),
Let \((X_i)_{i \in \iota }\) be a family of measurable real random variables on a probability space \((\Omega , \mu )\) that are mutually independent, and let \(s\) be a finite set of indices. Fix a radius \(t_{\max }\), and suppose that for each \(i \in s\) there is a constant \(c_i\) such that \(e^{tX_i}\) is \(\mu \)-integrable and \(\mathbf{E}_\mu [e^{tX_i}] \le e^{c_i t^2}\) for every \(t\) with \(|t| \le t_{\max }\). Then the sum \(\sum _{i \in s} X_i\) satisfies the same two-parameter sub-exponential bound with radius \(t_{\max }\) and constant \(\sum _{i \in s} c_i\): for every \(t\) with \(|t| \le t_{\max }\), the moment \(e^{t \sum _{i \in s} X_i}\) is \(\mu \)-integrable and \(\mathbf{E}_\mu \! \left[e^{t \sum _{i \in s} X_i}\right] \le e^{(\sum _{i \in s} c_i)\, t^2}\).
Let \(X, Y\) be measurable maps from a probability space \((\Omega , \mu )\) taking values in measurable spaces \(A\) and \(B\), and let \(X', Y'\) be measurable maps from a probability space \((\Omega ', \mu ')\) into the same spaces \(A\) and \(B\), with \(Y\) and \(Y'\) measurable. Suppose the pairs \((X, Y)\) and \((X', Y')\) are identically distributed, that is, they induce the same law on \(A \times B\) under \(\mu \) and \(\mu '\) respectively. Then for every measurable set \(s \subseteq B\), the maps \(X\) and \(X'\) are identically distributed with respect to the conditioned measures \(\mu [\, \cdot \mid Y^{-1}(s)]\) and \(\mu '[\, \cdot \mid Y'^{-1}(s)]\).
Let \(X,Y,Z\) be measurable random variables on a probability space \((\Omega ,\mu )\) and let \(X',Y',Z'\) be measurable random variables on a probability space \((\Omega ',\mu ')\), all taking values in countable measurable spaces and each assuming only finitely many values. If the triples \((X,Y,Z)\) and \((X',Y',Z')\) are identically distributed, meaning they induce the same joint law on \(S\times T\times U\), then their conditional mutual informations agree:
Let \((\Omega ,\mu )\) and \((\Omega ',\mu ')\) be finite probability spaces, and let \(S\), \(T\), \(U\) be finite sets. Consider random variables \(X\colon \Omega \to S\), \(Y\colon \Omega \to T\), \(Z\colon \Omega \to U\) defined on the first space and \(X'\colon \Omega '\to S\), \(Y'\colon \Omega '\to T\), \(Z'\colon \Omega '\to U\) defined on the second. If the joint random vectors \((X,Y,Z)\) and \((X',Y',Z')\) are identically distributed — that is, they induce the same probability measure on \(S\times T\times U\) — then their conditional mutual informations agree:
Let \(\mu \) be a finite measure on a measurable space \(\Omega \), and let \(S\) and \(T\) be finite sets. Let \(X : \Omega \to S\) and \(Y : \Omega \to T\) be measurable maps such that, for every \(x \in S\) and \(y \in T\),
where each measure is taken as a real number. Then \(X\) and \(Y\) are independent under \(\mu \).
Let \(X\), \(Y\), and \(Z\) be random variables of finite range on a probability space \((\Omega , \mu )\), taking values in countable measurable spaces, and let \(A = Z^{-1}(B)\) be the event that \(Z\) lands in a measurable set \(B\). Writing \(I[X : Y \mid Z;\, \nu ]\) for the conditional mutual information of \(X\) and \(Y\) given \(Z\) computed under a measure \(\nu \), and \(\mu [\, \cdot \mid A]\) for \(\mu \) conditioned on the event \(A\), one has
Let \(\mu \) be a finite measure on a measurable space \(\Omega \), let \(A\) be a measurable set, and let \(F \subseteq A\). Writing \(\mu (\, \cdot \mid A)\) for the conditional measure obtained by restricting \(\mu \) to \(A\) and normalising by \(\mu (A)\), and taking all masses as real numbers, one has
Let \(X\colon \Omega \to S\) and \(Y\colon \Omega \to T\) be measurable random variables on a probability space \((\Omega ,\mu )\) (allowing also \(\mu \) to be the zero measure), where \(S\) and \(T\) are countable and each of \(X\) and \(Y\) takes only finitely many values. If \(f\colon T\to V\) is a measurable injection into a countable space \(V\), then passing \(Y\) through \(f\) leaves the mutual information with \(X\) unchanged:
Let \((\Omega , \mu )\) be a measured space and let \(S\) and \(T\) be countable measurable spaces. Suppose \(X, X' \colon \Omega \to S\) and \(Y, Y' \colon \Omega \to T\) are maps with \(X\) and \(Y\) measurable, and that \(X = X'\) and \(Y = Y'\) hold \(\mu \)-almost everywhere. Then the mutual information is unchanged, \(I[X : Y \, ;\, \mu ] = I[X' : Y' \, ;\, \mu ]\).
Let \(\mu \) be a probability measure (or the zero measure) on a sample space \(\Omega \), and let \(X\colon \Omega \to S\) and \(Y\colon \Omega \to T\) be measurable random variables taking countably many values, each with finite range. If there is a constant \(c\in S\) such that \(X=c\) holds \(\mu \)-almost everywhere, then the mutual information between \(X\) and \(Y\) vanishes: \(I[X:Y;\mu ]=0\).
Let \(X \colon \Omega \to S\) and \(Y \colon \Omega \to T\) be measurable random variables taking values in countable spaces, defined on a space \(\Omega \) carrying a probability measure \(\mu \) (or the zero measure), and suppose each of \(X\) and \(Y\) takes only finitely many values. If \(Y\) is \(\mu \)-almost-surely equal to a fixed constant \(c \in T\), then the mutual information between \(X\) and \(Y\) vanishes:
Let \(X\) and \(Y\) be measurable random variables on a measure space \((\Omega ,\mu )\), taking values in countable measurable spaces and each attaining only finitely many values, and suppose \(\mu \) is either the zero measure or a probability measure. Then the mutual information between \(X\) and \(Y\) is at most the Shannon entropy of \(X\):
Let \(X \colon \Omega \to S\) and \(Y \colon \Omega \to T\) be measurable random variables, each taking only finitely many values, defined on a space \(\Omega \) equipped with a measure \(\mu \) that is either the zero measure or a probability measure. Writing \(H[Y;\mu ]\) for the Shannon entropy of \(Y\) and \(I[X:Y;\mu ]\) for the mutual information between \(X\) and \(Y\), one has
Let \(X\), \(Y\), and \(W\) be measurable random variables on a probability space \((\Omega , \mu )\), taking values in countable measurable spaces, and suppose each of \(X\), \(Y\), \(W\) takes only finitely many values almost surely. Then the mutual information between \(X\) and the pair \((Y, W)\) decomposes as
where \(I(X : W \mid Y)\) is the mutual information of \(X\) and \(W\) conditioned on \(Y\). (If \(\mu \) is the zero measure rather than a probability measure, both sides are \(0\).)
Let \(\mu \) and \(\nu \) be probability measures on the two-point space \(\mathrm{Bool} = \{ \mathrm{false}, \mathrm{true}\} \), and suppose \(\nu \) has full support, meaning \(\nu (\{ b\} ) \gt 0\) for each of the two points \(b\). Writing \(D_{\mathrm{KL}}(\mu \, \| \, \nu ) \in [0, \infty ]\) for the Kullback–Leibler divergence valued in the extended nonnegative reals, its real cast agrees with the real-valued Kullback–Leibler divergence of the identity map, regarded as a random variable, under \(\mu \) and under \(\nu \); that is,
Let \(\mu \) be a finite measure on a measurable space \(\Omega \), let \(X:\Omega \to \{ 0,1\} \) be a measurable Boolean random variable, and let \(S\subseteq \Omega \) be a set whose \(\mu \)-measure (as a real number) is nonzero, so that the conditional probability measure \(\mu (\cdot \mid S)\) is well defined. Let \(\nu \) be a probability measure on \(\{ 0,1\} \) that assigns strictly positive mass to each of the two points. Then the extended-real Kullback–Leibler divergence of the law \(X_{\ast }\, \mu (\cdot \mid S)\) of \(X\) under the conditional measure relative to \(\nu \), taken as a real number, coincides with its explicit pointwise value:
Let \(p\) be a mixed strategy over \(M\) actions for the row player and \(q\) a mixed strategy over \(N\) actions for the column player, and let \(\sigma \) be their product distribution, which assigns probability \(p_i\, q_j\) to each pure profile \((i,j)\). Then for every column action \(j\), the column marginal of \(\sigma \) at \(j\) equals \(q_j\); that is, \(\sum _{i} p_i\, q_j = q_j\).
Let \(p\) be a mixed strategy for the row player over \(M\) actions and \(q\) a mixed strategy for the column player over \(N\) actions, and let \(\sigma \) be their product distribution, so that \(\sigma _{ij} = p_i q_j\) for every profile \((i,j)\). Then for each row action \(i\) the row marginal of \(\sigma \) equals the corresponding weight of \(p\), that is, \(\sum _{j} p_i q_j = p_i\).
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N \ge 1\) column actions, and payoff matrix \(A\). Fix a row action \(i\) and a sequence of column actions \(a_0, \dots , a_{T-1}\) with \(T \gt 0\), and let \(\hat{q}\) be the associated empirical column strategy. Then the expected payoff of row \(i\) against \(\hat{q}\) equals the average payoff over the sequence:
Let \(C \le \mathcal{H}_n\) be a subspace of the \(n\)-qubit Hilbert space whose complex dimension is \(\dim _{\mathbb {C}} C = 2^k\), and suppose \(k \le n\). If \(C\) is a non-degenerate code correcting \(t\) errors — meaning it satisfies the Knill–Laflamme condition \(P_C\, E^\dagger F\, P_C = \lambda _{EF}\, P_C\) for all Pauli strings \(E,F\) of weight at most \(t\), and moreover \(P_C\, E^\dagger F\, P_C = 0\) whenever \(E \ne F\), where \(P_C\) is the orthogonal projection onto \(C\) — then
Let \(C \le \mathcal{H}_n\) be a subspace of the \(n\)-qubit Hilbert space that is a non-degenerate quantum code correcting \(t\) errors, and suppose its complex dimension is \(\dim _{\mathbb {C}} C = 2^{k}\) for some natural number \(k\). Then
Let \(p\) be prime and work in the symplectic space \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) equipped with the form \(\omega \). If \(S \le V\) is an isotropic subspace, so that \(\omega (u,v) = 0\) for all \(u, v \in S\), then its logical dimension \(k(S) = n - \dim (S)\) and its distance \(d(S)\) satisfy
Let \(\mathrm{Rad} = \tfrac {1}{2}\delta _{-1} + \tfrac {1}{2}\delta _{+1}\) be the Rademacher distribution on \(\mathbb {R}\), and let \(\mu \) be the product measure on the space of \(k \times d\) real arrays \(\omega = (\omega _{ij})\) under which all \(k \cdot d\) entries are independent and \(\mathrm{Rad}\)-distributed. Then for every index \((i,j) \in \mathrm{Fin}\, k \times \mathrm{Fin}\, d\), the pushforward of \(\mu \) along the coordinate projection \(\omega \mapsto \omega _{ij}\) is exactly \(\mathrm{Rad}\).
Draw a \(k \times d\) real matrix \(\omega = (\omega _{ij})\) from the joint measure under which all \(k \cdot d\) entries are independent and identically distributed Rademacher variables, each taking the values \(-1\) and \(+1\) with probability \(1/2\). Then for every fixed row index \(i \in \{ 1,\dots ,k\} \), the family of coordinate maps \(\bigl(\omega \mapsto \omega _{ij}\bigr)_{1 \le j \le d}\) — the \(d\) entries lying in row \(i\) — is mutually independent.
Fix natural numbers \(k\) and \(d\) and a row index \(i \in \{ 0, \dots , k-1\} \), and consider the joint measure on the set of \(k \times d\) real arrays under which all \(k \cdot d\) entries are independent Rademacher variables, each taking the values \(-1\) and \(+1\) with probability \(\tfrac 12\) apiece. Then the pushforward of this measure along the map that sends an array to its \(i\)-th row, an element of \(\mathbb {R}^d\), equals the \(d\)-fold product \(\bigotimes _{j=1}^{d} \mathrm{Rad}\) of Rademacher distributions.
Equip the space of \(k \times d\) real arrays \(\omega = (\omega _{ij})\) with the joint Rademacher product measure, under which the \(k \cdot d\) entries are independent and each \(\omega _{ij}\) takes the values \(-1\) and \(+1\) with equal probability \(1/2\). For each index \(i \in \{ 1,\dots ,k\} \), let \(R_i\) be the \(i\)-th row, that is, the \(\mathbb {R}^d\)-valued random variable \(\omega \mapsto (\omega _{ij})_{j}\). Then the family \((R_i)_{i}\) of \(k\) rows is mutually independent.
Fix natural numbers \(k\) and \(d\), and let the entries of a \(k \times d\) matrix be drawn independently from the Rademacher distribution and each scaled by \(1/\sqrt{k}\), giving a random matrix \(A(\omega ) \in \mathbb {R}^{k \times d}\) whose sample \(\omega \) is distributed according to the joint Rademacher product measure. Then for every fixed vector \(x \in \mathbb {R}^{d}\), the \(k\) real-valued random variables \(\omega \mapsto \bigl(A(\omega )\, x\bigr)_i\), one for each row index \(i \in \{ 1,\dots ,k\} \), are mutually independent.
Fix natural numbers \(k, d\), a vector \(x \in \mathbb {R}^d\), and an index \(i \in \{ 1, \dots , k\} \). Regard the sample space \(\mathrm{Rad\Omega }(k,d)\) of \(k \times d\) real arrays with its product measurable structure, and for a sample \(\omega \) let \(A(\omega )\) be the Rademacher matrix with entries \(A(\omega )_{ij} = \omega _{ij}/\sqrt{k}\), acting as a linear map \(\mathbb {R}^d \to \mathbb {R}^k\). Then the map \(\omega \mapsto (A(\omega )\, x)_i\), sending each sample to the \(i\)-th coordinate of the image of \(x\), is measurable.
Fix natural numbers \(k, d\), a vector \(x \in \mathbb {R}^d\), and a row index \(i \in \{ 1,\dots ,k\} \). For a sample \(\omega \) in the Rademacher sample space \(\mathrm{Rad\Omega }(k,d)\), let \(A(\omega )\) be the Rademacher matrix with entries \(\omega _{ij}/\sqrt{k}\). Then, as functions of \(\omega \), the \(i\)-th coordinate of the image \(A(\omega )x\) equals the sum of the scaled entries:
Fix natural numbers \(k, d\), a vector \(x \in \mathbb {R}^d\), and a row index \(i \in \{ 1,\dots ,k\} \). Under the joint Rademacher product measure on \(\mathrm{Rad\Omega }(k,d)\), the family of real random variables
indexed by the columns of row \(i\), is mutually independent.
Let \(p\) be a prime, and let \(\mathrm{op}\) be a gate operation belonging to the \(\mathrm{AC}^0[p]\) gate set over the alphabet \(\mathrm{Fin}\, 2\). Then \(\mathrm{op}\) is one of exactly four kinds of gate: the identity gate \(x \mapsto x_0\) on a single input, the NOT gate \(x \mapsto 1 - x_0\) on a single input, the unbounded fan-in AND gate \(x \mapsto \prod _{i} x_i\) for some arity \(n\), or the unbounded \(\mathrm{MOD}_p\) gate for some arity \(n\).
Let \(K\) be a field, let \(\omega \in K\), let \(n\) be a natural number, and let \(i \in \{ 0,\dots ,n-1\} \) index one of the \(n\) coordinates. Then for every point \(x = (x_0,\dots ,x_{n-1}) \in K^n\), evaluating the affine inverse coordinate polynomial \(1 + \omega ^{-1} - \omega ^{-1} X_i\) at \(x\) yields
Let \(K\) be a field, fix an element \(\omega \in K\), and let \(n\) be a natural number with a distinguished coordinate index \(i \in \{ 0,\dots ,n-1\} \). The polynomial \(1 + \omega ^{-1} - \omega ^{-1} X_i\) in the polynomial ring \(K[X_0,\dots ,X_{n-1}]\) has total degree at most \(1\).
Let \(K\) be a field, let \(\omega \in K\), and fix \(n \in \mathbb {N}\). For a subset \(s \subseteq \{ 1,\dots ,n\} \), form the polynomial \(\prod _{i \in s}\bigl(1 + \omega ^{-1} - \omega ^{-1} X_i\bigr) \in K[X_1,\dots ,X_n]\). Then for every point \(x \colon \{ 1,\dots ,n\} \to K\), this polynomial evaluates at \(x\) to
Let \(K\) be a field, let \(\omega \in K\), and fix \(n \in \mathbb {N}\) together with a subset \(s \subseteq \{ 0, 1, \dots , n-1\} \). Then the total degree of the affine-substituted squarefree monomial \(\prod _{i \in s}\bigl(1 + \omega ^{-1} - \omega ^{-1} X_i\bigr)\), viewed as a polynomial in \(K[X_0, \dots , X_{n-1}]\), is at most \(\left\lvert s\right\rvert \).
Fix a prime \(p\), polynomials \(P_1,\dots ,P_{\mathrm{width}}\) over \(\mathbb {Z}/p\) in \(\mathrm{vars}\) variables, and an evaluation point \(y \in (\mathbb {Z}/p)^{\mathrm{vars}}\). Consider the \(2^{\mathrm{width}\cdot \ell }\) possible seeds \(S = (S_1,\dots ,S_\ell )\), each \(S_k\) a subset of \(\{ 1,\dots ,\mathrm{width}\} \), and let \(N\) be the number of seeds for which the randomized AND-approximating polynomial evaluated at \(y\) differs from \(\prod _{k=1}^{\mathrm{width}}\bigl(1-(1-P_k(y))^{p-1}\bigr)\). Then
Fix a prime \(p\), and let \(P_1,\dots ,P_{\mathrm{width}}\) be polynomials over \(\mathbb {Z}/p\) in the input variables, together with subsets \(S_1,\dots ,S_\ell \) of the index set \(\{ 1,\dots ,\mathrm{width}\} \). Then the AND-approximating polynomial built from these data has total degree at most
Fix a prime \(p\) and, for natural numbers \(\mathrm{vars}\), \(\mathrm{width}\), and \(\ell \), let \(P_1,\dots ,P_{\mathrm{width}}\) be polynomials in \(\mathrm{vars}\) variables over \(\mathbb {Z}/p\). A random seed is a tuple \((S_1,\dots ,S_\ell )\) of subsets of the width positions \(\{ 1,\dots ,\mathrm{width}\} \), of which there are \((2^{\mathrm{width}})^{\ell } = 2^{\mathrm{width}\cdot \ell }\). The list obtained by recording the AND-approximating polynomial of \(P_1,\dots ,P_{\mathrm{width}}\) for every such seed therefore has length \(2^{\mathrm{width}\cdot \ell }\).
Given polynomials \(P_1,\dots ,P_{\mathrm{width}}\) over \(\mathbb {Z}/p\) and a random seed consisting of \(\ell \) subsets \(S_1,\dots ,S_\ell \) of the input positions, the approximator is
Fix a prime \(p\), and let \(P_1,\dots ,P_{\mathrm{width}}\) be polynomials in the variables \(y_1,\dots ,y_{\mathrm{vars}}\) over \(\mathbb {Z}/p\), together with subsets \(S_1,\dots ,S_\ell \) of the index set \(\{ 1,\dots ,\mathrm{width}\} \). Then for every point \(y : \{ 1,\dots ,\mathrm{vars}\} \to \mathbb {Z}/p\), evaluating the randomized OR-approximating polynomial \(1 - \prod _{k=1}^{\ell }\bigl(1 - (\sum _{i \in S_k} P_i)^{p-1}\bigr)\) at \(y\) yields the value obtained by feeding the scalars \(P_1(y),\dots ,P_{\mathrm{width}}(y)\) into the value-level OR approximator with the same subsets:
Fix a prime \(p\), and let \(v : \mathrm{Fin}\, \mathrm{width} \to \mathbb {Z}/p\) be a vector together with a family \(S_1,\dots ,S_\ell \) of subsets of \(\{ 1,\dots ,\mathrm{width}\} \). Then the value-level OR approximator built from \(v\) and \(S\) disagrees with the value-level OR detector of \(v\) if and only if \(v \neq 0\) yet \(\sum _{i \in S_k} v_i = 0\) in \(\mathbb {Z}/p\) for every index \(k\).
Fix a prime \(p\), and let \(P_1,\dots ,P_{\mathrm{width}}\) be polynomials in \(\mathrm{vars}\) variables over \(\mathbb {Z}/p\), evaluated at a fixed point \(y \in (\mathbb {Z}/p)^{\mathrm{vars}}\). A seed is a tuple \(S = (S_1,\dots ,S_\ell )\) of \(\ell \) subsets of \(\{ 1,\dots ,\mathrm{width}\} \), of which there are \(2^{\mathrm{width}\cdot \ell }\) in all; call \(S\) bad at \(y\) when the randomized OR-approximating polynomial built from the \(P_i\) and \(S\), evaluated at \(y\), differs from the value-level OR
Then the number of seeds that are bad at \(y\), multiplied by \(2^\ell \), is at most \(2^{\mathrm{width}\cdot \ell }\); equivalently, at most a \(2^{-\ell }\) fraction of all seeds are bad at \(y\).
Let \(p\) be a prime, and work over the ring of polynomials in \(\mathrm{vars}\) variables with coefficients in \(\mathbb {Z}/p\). Given a family of polynomials \(P_1,\dots ,P_{\mathrm{width}}\) from this ring, together with a seed consisting of \(\ell \) subsets \(S_1,\dots ,S_\ell \) of the index set \(\{ 1,\dots ,\mathrm{width}\} \), the associated OR-approximating polynomial has total degree at most
where the supremum is taken over all indices \(i\).
Fix a prime \(p\) and, for a given width \(\mathrm{width}\) and seed length \(\ell \), a family of polynomials \(P_1,\dots ,P_{\mathrm{width}}\) over \(\mathbb {Z}/p\) in a fixed number of variables. Form the list whose entries are the OR-approximating polynomials, one for each random seed — a seed being a tuple \((S_1,\dots ,S_\ell )\) of \(\ell \) subsets of the index set \(\{ 1,\dots ,\mathrm{width}\} \), listed with multiplicity. Then this list has length \(2^{\mathrm{width}\cdot \ell }\), the number of such seeds.
For natural numbers \(\mathrm{width}\) and \(\ell \), the number of \(\ell \)-tuples of subsets of an \(\mathrm{width}\)-element set equals \(2^{\mathrm{width}\cdot \ell }\); that is, the set of functions assigning to each index \(i \in \{ 0,\dots ,\ell -1\} \) a subset of \(\{ 0,\dots ,\mathrm{width}-1\} \) has exactly \(2^{\mathrm{width}\cdot \ell }\) elements.
Let \(p\) be a prime and let \(a \in \mathbb {Z}/p\) lie in \(\{ 0, 1\} \). Then
where the right-hand side denotes the element of \(\{ 0,1\} \subseteq \mathbb {Z}/p\) obtained by casting the value \(\mathrm{bitify}(a) \in \mathrm{Fin}\, 2\) through the natural numbers into \(\mathbb {Z}/p\); that is, the expression \(1 - (1 - a)^{p-1}\) equals \(1\) when \(a = 1\) and \(0\) when \(a = 0\).
Let \(p\) be a prime and let \(b \in \mathrm{Fin}\, 2\) be a bit. Booleanizing the field value of \(b\) recovers \(b\): taking the element of \(\mathbb {Z}/p\) represented by \(b\) and then mapping it back to \(\mathrm{Fin}\, 2\) (sending \(1\) to \(1\) and everything else to \(0\)) yields \(b\) again.
Let \(K\) be a field and let \(q\) be a positive natural number. For a residue \(r \in \{ 0,\dots ,q-1\} \), the Boolean residue-\(r\) indicator is the \(K\)-valued function on Boolean inputs \(x \in \{ 0,1\} ^n\) that is \(1\) when the Hamming weight of \(x\) is congruent to \(r\) modulo \(q\), and \(0\) otherwise.
Let \(K\) be a field, \(\omega \in K\), and let \(x \in \{ 0,1\} ^n\) be a Boolean input. The product of the coordinates of the point of \(\{ 1,\omega \} ^n\) obtained from \(x\) by the embedding \(0 \mapsto 1\), \(1 \mapsto \omega \) equals \(\omega ^{\left\lvert x\right\rvert }\), where \(\left\lvert x\right\rvert = \sum _i x_i\) is the Hamming weight of \(x\).
Let \(p\) be a prime and let \(b : \mathrm{Fin}\, 2\) be a bit. Then the element of \(\mathbb {Z}/p\) represented by \(b\) — the image of the natural number \(b\) under the canonical map \(\mathbb {N}\to \mathbb {Z}/p\) — belongs to the subset \(\{ 0,1\} \) of \(\mathbb {Z}/p\).
Let \(p\) be a prime and let \(a \in \mathbb {Z}/p\) satisfy \(a \in \{ 0, 1\} \). Write \(\beta (a) \in \mathrm{Fin}\, 2\) for the booleanization of \(a\), equal to \(1\) when \(a = 1\) and to \(0\) otherwise. Then interpreting \(\beta (a)\) as a natural number and reducing it back modulo \(p\) recovers \(a\); that is, the image of \(\beta (a)\) under \(\mathrm{Fin}\, 2 \to \mathbb {N} \to \mathbb {Z}/p\) equals \(a\).
Fix a prime \(p\) and let \(v : \mathrm{Fin}\, \mathrm{width} \to \mathbb {Z}/p\) be a nonzero vector. Consider families \(S = (S_1,\dots ,S_\ell )\) of subsets \(S_k \subseteq \mathrm{Fin}\, \mathrm{width}\), indexed by \(\mathrm{Fin}\, \ell \), and call such a family *bad* for \(v\) if the value-level OR approximator \(1 - \prod _k\bigl(1 - (\sum _{i \in S_k} v_i)^{p-1}\bigr)\) disagrees with the value-level OR detector \(1 - \prod _k\bigl(1 - v_k^{\, p-1}\bigr)\). Then the number of bad families, multiplied by \(2^\ell \), is at most the total number of families \(S : \mathrm{Fin}\, \ell \to \mathrm{Finset}(\mathrm{Fin}\, \mathrm{width})\).
Fix a prime \(p\) and natural numbers \(w\) and \(\ell \), and let \(v = (v_1,\dots ,v_w)\) be an arbitrary vector in \((\mathbb {Z}/p)^w\) (no coordinate is assumed nonzero). Call a *seed* a tuple \(S = (S_1,\dots ,S_\ell )\) of subsets of \(\{ 1,\dots ,w\} \), and call it *bad* when the value-level OR approximator determined by \(v\) and \(S\) differs from the value-level OR detector of \(v\). Then, writing \(N = (2^{w})^{\ell }\) for the total number of seeds, the number of bad seeds satisfies
Let \(F\) be a feedforward circuit over an alphabet \(\alpha \), let \(d\) be a layer index with \(0 \le d \lt \mathrm{depth}(F)\), let \(u\) be a node of layer \(d+1\), and let \(x\) be an assignment of alphabet values to the input nodes of \(F\). Writing \(g\) for the gate attached to \(u\), the value of \(u\) under \(x\) equals the gate operation of \(g\) applied to the tuple whose entry at each input slot \(i\) is the value under \(x\) of the layer-\(d\) node to which slot \(i\) is wired.
Let \(p\) be a prime, and let \(x_0,\dots ,x_{w-1}\) be elements of \(\mathbb {Z}/p\), each lying in \(\{ 0,1\} \). Then
where \(\beta \) is the booleanization \(\mathbb {Z}/p \to \mathrm{Fin}\, 2\) (sending \(1\) to \(1\) and everything else to \(0\)). The left-hand product is taken in \(\mathbb {Z}/p\); on the right, the product of the \(\beta (x_i)\) is formed in \(\mathrm{Fin}\, 2\), and its value—the natural number \(0\) or \(1\)—is cast into \(\mathbb {Z}/p\). Thus the Fermat-power expression on the left evaluates to the Boolean AND of the inputs.
Let \(p\) be a prime and let \(\mathrm{inputs}_0, \dots , \mathrm{inputs}_{w-1}\) be elements of \(\mathbb {Z}/p\), each lying in \(\{ 0,1\} \). Then
where the right-hand side is the output of the unbounded \(\mathrm{MOD}_p\) gate on the booleanizations of the inputs — equal to \(1\) when \(\sum _i \mathrm{inputs}_i = 0\) in \(\mathbb {Z}/p\) and to \(0\) otherwise — regarded as an element of \(\{ 0,1\} \subseteq \mathbb {Z}/p\).
Let \(p\) be a prime and fix natural numbers \(\mathrm{vars}\) and \(\mathrm{width}\). Let \(P_0,\dots ,P_{\mathrm{width}-1}\) be polynomials in \(\mathrm{vars}\) variables over the field \(\mathbb {Z}/p\). Then the total degree of the exact \(\mathrm{MOD}_p\) polynomial \(1 - \bigl(\sum _i P_i\bigr)^{p-1}\) is at most
the supremum being taken over all \(i\) (and read as \(0\) when \(\mathrm{width}=0\)).
Let \(p\) be a prime, and fix natural numbers \(n\) and \(\ell \). For every gate operation \(\mathrm{op}\) belonging to the \(\mathrm{AC}^0[p]\) gate set — the identity, the NOT gate, the unbounded fan-in AND gate of any arity, and the unbounded \(\mathrm{MOD}_p\) gate of any arity — there exists a gate approximator family for \(\mathrm{op}\) in \(n\) variables with parameter \(\ell \); that is, a nonempty finite seed type together with, for each tuple of incoming polynomials over \(\mathbb {Z}/p\) and each seed, an approximating polynomial in \(n\) variables whose total degree is at most \((p-1)\, \ell \) times the largest degree among the incoming polynomials, and such that for every Boolean input at which all incoming polynomials evaluate into \(\{ 0,1\} \), the number of seeds on which the approximator disagrees with the true gate value, multiplied by \(2^{\ell }\), is at most the total number of seeds.
Let \(p\) be prime, let \(\mathrm{vars}\), \(\mathrm{width}\), and \(\ell \) be natural numbers, and let \(P_1,\dots ,P_{\mathrm{width}}\) be polynomials over \(\mathbb {Z}/p\) in \(\mathrm{vars}\) variables. Call a *seed* an \(\ell \)-tuple \(S=(S_1,\dots ,S_\ell )\) of subsets of \(\{ 1,\dots ,\mathrm{width}\} \), and for each seed \(S\) write \(A_S\) for the associated AND-approximating polynomial; there are exactly \(2^{\mathrm{width}\cdot \ell }\) seeds. Then the list whose entries are the polynomials \(A_S\), one per seed, has length \(2^{\mathrm{width}\cdot \ell }\), and every entry has total degree at most \((p-1)\, \ell \, \sup _i \deg (P_i)\). Moreover, for every point \(y\in (\mathbb {Z}/p)^{\mathrm{vars}}\), if \(N(y)\) denotes the number of seeds \(S\) for which
(the right-hand side being the exact conjunction value at \(y\)), then \(N(y)\cdot 2^{\ell }\le 2^{\mathrm{width}\cdot \ell }\).
Fix a prime \(p\) and natural numbers \(\mathrm{vars}\), \(\mathrm{width}\), and \(\ell \), and let \(P_1,\dots ,P_{\mathrm{width}}\) be polynomials in \(\mathrm{vars}\) variables over \(\mathbb {Z}/p\). Then there is a list \(Ps\) of polynomials over \(\mathbb {Z}/p\) that is exactly the list of OR-approximating polynomials — one entry for every seed \(S\), a seed being an assignment \(k \mapsto S_k\) of a subset \(S_k\) of the \(\mathrm{width}\) positions to each of the \(\ell \) rounds — and this list has the following three properties. Its length is \(2^{\mathrm{width}\cdot \ell }\); every entry has total degree at most \((p-1)\, \ell \, \sup _i \deg (P_i)\); and for every evaluation point \(y \in (\mathbb {Z}/p)^{\mathrm{vars}}\), the number of seeds \(S\) for which the OR-approximating polynomial of \(S\) evaluated at \(y\) differs from the exact value \(1 - \prod _{k=1}^{\mathrm{width}}\bigl(1 - P_k(y)^{p-1}\bigr)\), when multiplied by \(2^{\ell }\), is at most the length \(2^{\mathrm{width}\cdot \ell }\) of \(Ps\).
Let \(\alpha \) and \(\beta \) be finite types with \(\beta \) nonempty, let \(\mathrm{Fail} : \alpha \to \beta \to \mathrm{Prop}\) be a decidable binary relation, and let \(B, C\) be natural numbers. Suppose that for every \(a \in \alpha \) the number of \(b \in \beta \) with \(\mathrm{Fail}\, a\, b\) satisfies
Then there exists a single \(b \in \beta \) for which
Let \(K\) be a field, let \(\omega \in K\), and let \(n\) be a natural number. Write \(\{ 1,\omega \} ^n\) for the root cube, the set of vectors \(x = (x_1,\dots ,x_n) \in K^n\) with \(x_i \in \{ 1,\omega \} \) for every \(i\). Then for every function \(f : \{ 1,\omega \} ^n \to K\) there exists a polynomial \(P \in K[X_1,\dots ,X_n]\) such that \(P(x_1,\dots ,x_n) = f(x)\) for all \(x\) in the cube.
Let \(p\) be prime and let \(n, \ell \in \mathbb {N}\). Let \(F\) be a feedforward circuit over the alphabet \(\mathrm{Fin}\, 2\) with input type \(\mathrm{Fin}\, n\), a single output node, and a finite node set at every layer, each of whose gates is an \(\mathrm{AC}^0[p]\) gate (the identity, NOT, an unbounded fan-in AND, or an unbounded \(\mathrm{MOD}_p\) gate of some arity); write \(d\) for its depth. For a Boolean input \(x : \mathrm{Fin}\, n \to \mathrm{Fin}\, 2\), let \(\hat{x}\) be its coordinatewise image in \(\mathbb {Z}/p\) and let \(F(x) \in \{ 0,1\} \subseteq \mathbb {Z}/p\) be the resulting single output value cast into the field. Then there exist a nonempty finite seed set \(\mathrm{Seed}\) and a family of polynomials \(P_s \in (\mathbb {Z}/p)[X_0,\dots ,X_{n-1}]\) indexed by \(s \in \mathrm{Seed}\), each of total degree at most \(\bigl((p-1)\ell \bigr)^{d}\), such that for every Boolean input \(x\),
where \(G\) is the total number of non-input gates of \(F\), that is, the total number of nodes across layers \(1,\dots ,d\).
Let \(p\) be a prime, and let \(F\) be a feedforward circuit over the Boolean alphabet \(\mathrm{Fin}\, 2\) with input coordinates indexed by \(\mathrm{Fin}\, n\) and a single output node, in which every layer has finitely many nodes and every gate operation belongs to the \(\mathrm{AC}^0[p]\) gate set (the identity gate, the NOT gate, the unbounded fan-in AND gate of every arity, and the \(\mathrm{MOD}_p\) gate of every arity). For a Boolean input \(x : \mathrm{Fin}\, n \to \mathrm{Fin}\, 2\), write \(\bar{x} \in (\mathbb {Z}/p)^n\) for the point whose \(i\)-th coordinate is the residue of \(x_i\) modulo \(p\), and let \(c(x) \in \mathbb {Z}/p\) be the residue of the single Boolean value that \(F\) computes on \(x\). Then for every \(\ell \in \mathbb {N}\) there exist a finite nonempty index set \(\mathrm{Seed}\) and a family of polynomials \((P_s)_{s \in \mathrm{Seed}}\) in \(\mathbb {Z}/p[X_1,\dots ,X_n]\), each of total degree at most \(\bigl((p-1)\ell \bigr)^{d}\), where \(d\) is the depth of \(F\), such that for every Boolean input \(x\),
where \(\operatorname {size}(F)\) denotes the number of non-input nodes of \(F\).
Let \(p\) be a prime, let \(F\) be a feedforward circuit over the alphabet \(\mathrm{Fin}\, 2\) with input type \(\mathrm{Fin}\, n\) and a finite output type, all of whose node sets are finite and which uses only \(\mathrm{AC}^0[p]\) gates, and fix an error parameter \(\ell \in \mathbb {N}\). Write \(d = \mathrm{depth}(F)\) for its depth and \(N\) for the total number of its non-input gates, that is, the sum over the layers \(1,\dots ,d\) of the number of nodes in each layer. For a bit \(b \in \mathrm{Fin}\, 2\) write \(\overline{b} \in \mathbb {Z}/p\) for the image of \(b\) regarded as a natural number, and for a Boolean input \(x : \mathrm{Fin}\, n \to \mathrm{Fin}\, 2\) write \(\hat{x} \in (\mathbb {Z}/p)^n\) for the point whose \(i\)-th coordinate is \(\overline{x_i}\). Then there exist a nonempty finite seed type \(S\) and, for each seed \(s \in S\) and each output node \(o\), a polynomial \(P_{s,o}\) in \(n\) variables over \(\mathbb {Z}/p\), such that every \(P_{s,o}\) has total degree at most \(\bigl((p-1)\ell \bigr)^{d}\) and, for every Boolean input \(x\),
where \(F(x)(o) \in \mathrm{Fin}\, 2\) is the value the circuit \(F\) computes at output node \(o\) on input \(x\).
Let \(p\) be a prime and let \(F\) be a feedforward circuit over the Boolean alphabet \(\{ 0,1\} \) with inputs indexed by \(\{ 1,\dots ,n\} \) and a finite set of outputs, in which every layer has only finitely many nodes and every gate belongs to the \(\mathrm{AC}^0[p]\) gate set. Then for every \(\ell \in \mathbb {N}\) there exist a finite, nonempty seed set \(\mathrm{Seed}\) and, for each seed \(s \in \mathrm{Seed}\) and each output \(o\), a polynomial \(P_{s,o} \in \mathbb {Z}/p[X_1,\dots ,X_n]\) of total degree at most \(\bigl((p-1)\ell \bigr)^{\mathrm{depth}(F)}\), such that for every Boolean input \(x \in \{ 0,1\} ^n\) the seeds that err on \(x\) satisfy
where \(\bar x \in (\mathbb {Z}/p)^n\) and \(\overline{F(x)(o)} \in \mathbb {Z}/p\) denote the images under \(\{ 0,1\} \to \mathbb {Z}/p\) of the Boolean vector \(x\) and of the circuit’s output at \(o\) on input \(x\), respectively.
Let \(p\) be a prime, and fix natural numbers \(n\) and \(\ell \). Let \(\mathrm{op}\) be a gate operation from the \(\mathrm{AC}^0[p]\) gate set, and for each input wire \(i\) of \(\mathrm{op}\) let \(P_i\) be a polynomial in \(\mathbb {Z}/p[X_1,\dots ,X_n]\). Then there exist a nonempty finite index set \(\mathrm{Seed}\) with decidable equality and a family of polynomials \((P_s)_{s\in \mathrm{Seed}}\) in \(\mathbb {Z}/p[X_1,\dots ,X_n]\) with the following two properties. First, each \(P_s\) has total degree at most \((p-1)\, \ell \, \sup _i \deg (P_i)\). Second, for every Boolean assignment \(x\in \{ 0,1\} ^n\), write \(y\in (\mathbb {Z}/p)^n\) for the coordinatewise image of \(x\) under the inclusion \(\mathrm{Fin}\, 2\hookrightarrow \mathbb {Z}/p\); then, provided every value \(P_i(y)\) lies in \(\{ 0,1\} \), the number of seeds \(s\in \mathrm{Seed}\) for which \(P_s(y)\) differs from the gate output \(\mathrm{op}\bigl(i\mapsto \mathrm{bitify}(P_i(y))\bigr)\), multiplied by \(2^\ell \), is at most \(\left\lvert \mathrm{Seed}\right\rvert \).
Let \(p\) be a prime and let \(F\) be a layered feedforward circuit over the Boolean alphabet \(\mathrm{Fin}\, 2\), with inputs indexed by \(\mathrm{Fin}\, n\), finitely many nodes in every layer, and a single output node, and suppose every gate of \(F\) belongs to the \(\mathrm{AC}^0[p]\) gate set (the identity, NOT, unbounded fan-in AND, and unbounded \(\mathrm{MOD}_p\) gates of every arity). Then for every \(\ell \in \mathbb {N}\) there is a nonempty list \(P_1,\dots ,P_m\) of polynomials in the variables \(\{ X_i : i \in \mathrm{Fin}\, n\} \) over \(\mathbb {Z}/p\), each of total degree at most \(\bigl((p-1)\ell \bigr)^{\mathrm{depth}(F)}\), with the following property. For a Boolean input \(x : \mathrm{Fin}\, n \to \mathrm{Fin}\, 2\), write \(\hat{x} \in (\mathbb {Z}/p)^n\) for the point whose \(i\)-th coordinate is the image of \(x_i\) in \(\mathbb {Z}/p\), and let \(b(x) \in \mathbb {Z}/p\) be the image of the single Boolean value output by \(F\) on \(x\). Then for every such \(x\), the number of indices \(j\) with \(P_j(\hat{x}) \neq b(x)\), multiplied by \(2^{\ell }\), is at most \(s \cdot m\), where \(s\) is the total number of non-input gates of \(F\) and \(m\) is the length of the list.
Let \(p\) be a prime and let \(F\) be a feedforward circuit over the alphabet \(\mathrm{Fin}\, 2\) with \(n\) Boolean input variables, finitely many nodes in each layer, and a single output node, every gate of which belongs to the \(\mathrm{AC}^0[p]\) gate set. Then for each \(\ell \in \mathbb {N}\) there exists a nonempty list of polynomials \(P_1, \dots , P_m \in (\mathbb {Z}/p)[X_1, \dots , X_n]\), each of total degree at most \(\bigl((p-1)\ell \bigr)^{\mathrm{depth}(F)}\), with the following property. Given a Boolean input \(x \colon \mathrm{Fin}\, n \to \mathrm{Fin}\, 2\), let \(\hat{x} \colon \mathrm{Fin}\, n \to \mathbb {Z}/p\) denote \(x\) cast coordinatewise into \(\mathbb {Z}/p\), and let \(b(x) \in \mathbb {Z}/p\) denote the image of the single output bit that \(F\) computes on \(x\). Then for every such \(x\) the number of indices \(j\) at which the polynomial disagrees with the circuit satisfies
where \(\mathrm{size}(F)\) is the number of non-input nodes of \(F\) and \(m\) is the length of the list.
Let \(p\) be a prime, let \(n\in \mathbb {N}\), and let \(F\) be a feedforward circuit over the alphabet \(\{ 0,1\} \) with input type \(\mathrm{Fin}\, n\), a unique output node, and finitely many nodes in each layer, every one of whose gates is drawn from the \(\mathrm{AC}^0[p]\) gate set (the identity, NOT, and unbounded fan-in AND gates, together with the unbounded \(\mathrm{MOD}_p\) gates). Write \(g\colon \{ 0,1\} ^n\to \mathbb {Z}/p\) for the function computed by \(F\), namely its single output value cast into \(\mathbb {Z}/p\). Then for every \(\ell \in \mathbb {N}\) there exists a polynomial \(P\in (\mathbb {Z}/p)[X_1,\dots ,X_n]\) of total degree at most \(\bigl((p-1)\ell \bigr)^{\operatorname {depth}(F)}\) such that the number \(b\) of Boolean inputs \(x\in \{ 0,1\} ^n\) at which \(P\), evaluated at the point of \((\mathbb {Z}/p)^n\) whose \(i\)-th coordinate is the image of \(x_i\) in \(\mathbb {Z}/p\), differs from \(g(x)\) satisfies
where \(\operatorname {size}(F)\) is the number of non-input nodes of \(F\).
Fix a prime \(p\), and let \((P_s)_{s \in S}\) be a family of polynomials in \((\mathbb {Z}/p)[X_1,\dots ,X_n]\) indexed by a nonempty finite set \(S\) of seeds, let \(f : \{ 0,1\} ^n \to \mathbb {Z}/p\) be a target function, and let \(\ell \) and \(B\) be natural numbers. For a Boolean input \(x \in \{ 0,1\} ^n\), write \(\bar{x} \in (\mathbb {Z}/p)^n\) for its coordinatewise image under the inclusion \(\{ 0,1\} \hookrightarrow \mathbb {Z}/p\). Suppose that for every Boolean input \(x\),
Then there exists a seed \(s \in S\) for which the number of Boolean inputs \(x \in \{ 0,1\} ^n\) with \(P_s(\bar{x}) \neq f(x)\), multiplied by \(2^{\ell }\), is at most \(B \cdot 2^{n}\).
Let \(K\) be a field, fix an element \(\omega \in K\), and let \(n\) be a natural number. Write \(\{ 1,\omega \} ^n\) for the set of vectors \(x = (x_1,\dots ,x_n)\) with each coordinate \(x_i \in \{ 1,\omega \} \). Then for every function \(f : \{ 1,\omega \} ^n \to K\) there exist coefficients \(c_S \in K\), indexed by the subsets \(S \subseteq \{ 1,\dots ,n\} \), such that the squarefree (multilinear) polynomial
satisfies \(P(x) = f(x)\) for every \(x \in \{ 1,\omega \} ^n\).
Let \(p\) and \(q\) be distinct prime numbers, and let \(\mathbb {F}_{p^{q-1}}\) denote the finite field with \(p^{q-1}\) elements. Then the multiplicative group \(\mathbb {F}_{p^{q-1}}^{\times }\) contains a unit \(u\) whose order is exactly \(q\).
Let \(\alpha \), \(\beta \), and \(\mathrm{Cand}\) be finite sets, and assign to each candidate \(c \in \mathrm{Cand}\) a function \(\mathrm{center}(c) \colon \alpha \to \beta \). For a radius \(e \in \mathbb {N}\), write the Hamming ball of radius \(e\) about \(c\) for the set of functions \(f \colon \alpha \to \beta \) that differ from \(\mathrm{center}(c)\) in at most \(e\) coordinates, that is, with \(\left\lvert \{ a \in \alpha : \mathrm{center}(c)(a) \neq f(a)\} \right\rvert \le e\). Suppose that every such ball contains at most \(B\) functions, and that every \(f \colon \alpha \to \beta \) lies in the ball about at least one candidate. Then
Let \(s\) be a finite set of elements of a type \(\alpha \), and let \(q\) be a decidable predicate on \(\alpha \). Writing the elements of \(s\) out as a list (in any order) and discarding those that fail \(q\) produces a list whose length equals the number of elements of \(s\) that satisfy \(q\); that is, the length of the filtered list agrees with the cardinality of the subset \(\{ x \in s : q(x)\} \).
Let \(\alpha \) and \(\beta \) be finite types, fix a function \(\mathrm{center} : \alpha \to \beta \), and let \(e \in \mathbb {N}\). Then the number of functions \(f : \alpha \to \beta \) that disagree with \(\mathrm{center}\) on at most \(e\) points of \(\alpha \) — that is, for which \(\left\lvert \{ a \in \alpha : \mathrm{center}(a) \ne f(a)\} \right\rvert \le e\) — satisfies
For a circuit \(F\) with finite node sets, \(\mathrm{gateCountBefore}\, F\, d\) counts the non-input gates in layers \(1,\dots ,d\): it is \(0\) for \(d = 0\) and, recursively, the count for \(d\) plus the number of nodes at layer \(d+1\).
Let \(F\) be a feedforward circuit over the alphabet \(\mathbb {F}_2\) with input type \(\mathrm{Fin}\, n\), and suppose each of its node layers is finite. Counting the non-input nodes layer by layer through all of the layers \(1,\dots ,\mathrm{depth}(F)\), and accumulating these counts over the full depth of \(F\), yields exactly the size of \(F\), that is, its total number of non-input nodes.
Let \(F\) be a Boolean feed-forward circuit (alphabet \(\mathbb {F}_2\), input type \(\mathrm{Fin}\, n\)) all of whose node layers are finite, and write \(\mathrm{depth}(F)\) for its depth. For a natural number \(d\) with \(d \le \mathrm{depth}(F)\), the number of non-input gates in the first \(d\) layers equals the total number of nodes across layers \(1,\dots ,d\):
where \(L_i\) denotes the set of nodes in the \(i\)-th layer of \(F\).
Let \(F\) be a layered feedforward circuit over the Boolean alphabet \(\mathbb {F}_2\) with input type \(\mathrm{Fin}\, n\), in which every layer has finitely many nodes, and for each natural number \(k \le \mathrm{depth}(F)\) let \(g_F(k)\) denote the number of non-input gates lying in layers \(1,\dots ,k\) (so \(g_F(0) = 0\)). Then for every \(d\) with \(d + 1 \le \mathrm{depth}(F)\),
where \(\left\lvert N_{d+1}\right\rvert \) is the number of nodes in layer \(d+1\).
Let \(F\) be a layered feedforward circuit over the Boolean alphabet \(\{ 0,1\} \), whose input layer carries \(n\) nodes and each of whose layers has a finite node set. Then the number of non-input gates lying in the first \(0\) layers is zero, since the input layer contributes no gates.
For a feed-forward circuit \(F\), a bound \(d \le F.\mathrm{depth}\) and an index \(j : \mathrm{Fin}\, d\), the layer index \(\mathrm{gateLayerIdx}\, F\, h_d\, j \in \mathrm{Fin}(F.\mathrm{depth}+1)\) is \(j+1\), i.e. the \((j+1)\)-st node layer of \(F\). It packages the non-input layers \(1, \dots , d\) as indices into \(F.\mathrm{nodes}\) so that they can be summed over.
A LayerPolyFamily for a circuit \(F\) at layer \(d\) consists of a nonempty finite seed type and, for each seed, a polynomial for every node of layer \(d\), such that each polynomial has total degree at most \(\mathrm{circuitDegreeBound}(p,\ell ,d)\) and, for every Boolean input \(x\), the number of seeds for which some layer-\(d\) node is mispredicted, multiplied by \(2^{\ell }\), is at most \(\mathrm{gateCountBefore}\, F\, d\) times the number of seeds.
Let \(l\) be a finite list with entries in a set \(\alpha \), let \(f\colon \alpha \to \beta \) be a function, and let \(P\) be a decidable predicate on \(\beta \). Then applying \(f\) to every entry of \(l\) and retaining those whose image satisfies \(P\) yields a list of the same length as retaining those entries \(a\) of \(l\) for which \(P(f(a))\) holds; that is, the sublist \(\bigl(\mathrm{map}\, f\, (l)\bigr)\) filtered by \(P\) and the sublist of \(l\) filtered by the pulled-back predicate \(a \mapsto P(f(a))\) have equal length.
Let \(K\) be a field and let \(\omega \in K\) satisfy \(\omega \ne 1\). For natural numbers \(n\) and \(D\), write \(\{ 1,\omega \} ^n\) for the set of vectors \(x : \mathrm{Fin}\, n \to K\) each of whose coordinates equals \(1\) or \(\omega \), and let \(\mathcal{S}_{n,D}\) be the collection of finsets \(s \subseteq \mathrm{Fin}\, n\) with \(\left\lvert s\right\rvert \le D\). Then for every polynomial \(Q \in K[X_i : i \in \mathrm{Fin}\, n]\) of total degree at most \(D\), there exist coefficients \(c : \mathcal{S}_{n,D} \to K\) such that the squarefree polynomial \(\sum _{\left\lvert s\right\rvert \le D} c(s)\, \prod _{i \in s} X_i\) agrees with \(Q\) at every point of \(\{ 1,\omega \} ^n\).
Fix natural numbers \(n\) and \(D\), and write \(\mathcal{S}_{n,D}\) for the type of supports of squarefree monomials of degree at most \(D\) in \(n\) variables. For any finite type \(K_0\), the number of functions from \(\mathcal{S}_{n,D}\) to \(K_0\) is
Given coefficients \(c\) indexed by \(\texttt{RazborovSmolensky.LowDegreeSupport}\ n\ D\), the multivariate polynomial
i.e. the sum of the squarefree monomials \(\texttt{RazborovSmolensky.squarefreeMonomial}\ s\) weighted by \(c\), with the terms of support size \(\gt D\) set to \(0\).
Fix a field \(K\) and natural numbers \(n\) and \(D\), and let \(\mathcal{S}_{n,D}\) be the set of subsets \(s \subseteq \mathrm{Fin}\, n\) with \(\left\lvert s\right\rvert \le D\). For a coefficient function \(c \colon \mathcal{S}_{n,D} \to K\), write \(P(c) = \sum _{s \in \mathcal{S}_{n,D}} c(s)\, \prod _{i \in s} X_i\) for the associated squarefree polynomial in \(K[X_i : i \in \mathrm{Fin}\, n]\). Then for any two coefficient functions \(c_1, c_2 \colon \mathcal{S}_{n,D} \to K\) one has
where \(c_1 + c_2\) denotes the pointwise sum.
Fix a field \(K\) and natural numbers \(n, D\), and write \(\mathcal{S}_{n,D}\) for the type of subsets \(s \subseteq \mathrm{Fin}\, n\) with \(\left\lvert s\right\rvert \le D\). For a coefficient function \(c \colon \mathcal{S}_{n,D} \to K\), let \(P(c)\) denote the multivariate polynomial \(\sum _{s \in \mathcal{S}_{n,D}} c(s)\, \prod _{i \in s} X_i\) over \(K\) in the variables \(X_0, \dots , X_{n-1}\). Then for any finite index set \(S\) and any family of coefficient functions \((c_i)_{i \in S}\) with each \(c_i \colon \mathcal{S}_{n,D} \to K\), the polynomial built from the pointwise sum of coefficients equals the sum of the individual polynomials:
Let \(K\) be a field and fix \(n, D \in \mathbb {N}\). For any assignment of coefficients \(c \colon \mathcal{S}_{n,D} \to K\), indexed by the subsets \(s \subseteq \mathrm{Fin}\, n\) with \(\left\lvert s\right\rvert \le D\), the polynomial
has total degree at most \(D\).
Let \(K\) be a field and let \(n, D \in \mathbb {N}\). Applying the low-degree squarefree polynomial construction to the identically zero coefficient function on \(\mathcal{S}_{n,D}\) produces the zero element of the polynomial ring \(K[X_0, \dots , X_{n-1}]\); that is,
For \(n, D \in \mathbb {N}\), the type of supports of squarefree monomials of degree at most \(D\): the subtype of finsets \(s \subseteq \mathrm{Fin}\, n\) with \(\left\lvert s\right\rvert \le D\). In the displays below this type is written \(\mathcal{S}_{n,D}\).
For natural numbers \(n\) and \(D\), let \(\mathcal{S}_{n,D}\) denote the (finite) type of low-degree squarefree supports on \(\mathrm{Fin}\, n\), that is, the subsets \(s \subseteq \mathrm{Fin}\, n\) with \(\left\lvert s\right\rvert \le D\). Then the number of such supports satisfies
The map sending a low-degree support \(s\) to the pair consisting of its exact cardinality \(\left\lvert s\right\rvert \in \mathrm{Fin}(D+1)\) together with \(s\) itself, viewed as an element of \(\{ t : \mathrm{Finset}(\mathrm{Fin}\, n) \mid \left\lvert t\right\rvert = \left\lvert s\right\rvert \} \). It is used as an injection for the binomial-sum bound rather than as a full equivalence.
Let \(p\) and \(q\) be distinct primes, let \(\omega \in \mathbb {F}_{p^{q-1}}\) satisfy \(\omega ^q = 1\) and \(\omega \neq 1\), and let \(d, E, e\) be natural numbers with \(q \cdot E \le e\). Suppose no polynomial in \(n\) variables over \(\mathbb {F}_{p^{q-1}}\) of total degree at most \(d\) disagrees with the root product \(x \mapsto \prod _i x_i\) on at most \(e\) points of \(\{ 1,\omega \} ^n\). Then every polynomial in \(n + (q-1)\) variables over \(\mathbb {Z}/p\) of total degree at most \(d\) disagrees with the \(\mathrm{MOD}_q\) function (the indicator that the Hamming weight is divisible by \(q\)) on at least \(E\) Boolean inputs.
Let \(p\) and \(q\) be distinct primes and write \(K = \mathbb {F}_{p^{q-1}}\). Let \(F\) be a feed-forward Boolean circuit on \(n + (q-1)\) input bits with a single output and finitely many nodes in each layer, using only \(\mathrm{AC}^0[p]\) gate operations, and suppose that \(F\) computes the \(\mathrm{MOD}_q\) function: on every input, \(F\) outputs \(1\) exactly when the Hamming weight of the input is divisible by \(q\). Let \(\delta , \ell , e, B\) be natural numbers such that \(q \cdot \bigl(\delta \cdot 2^{\, n+(q-1)}\bigr) \le e\), such that \(\sum _{t=0}^{e} \binom {2^n}{t}\, \left\lvert K\right\rvert ^{t} \le B\), and such that \(\left\lvert K\right\rvert ^{2^n} \gt N \cdot B\), where \(N\) is the number of \(K\)-valued functions on the collection of subsets \(S \subseteq \{ 1,\dots ,n\} \) with \(\left\lvert S\right\rvert \le \lfloor n/2 \rfloor + \bigl((p-1)\ell \bigr)^{\mathrm{depth}(F)}\). Then the size of \(F\) (its number of non-input gates) satisfies \(\delta \cdot 2^{\ell } \le \mathrm{size}(F)\).
Let \(K\) be a field and let \(\omega \in K\) satisfy \(\omega \ne 1\). Fix \(n, D \in \mathbb {N}\), and consider a monomial \(a\, X^m\) in the variables \(X_1,\dots ,X_n\) over \(K\), given by an exponent vector \(m : \mathrm{Fin}\, n \to \mathbb {N}\) and a coefficient \(a \in K\), whose total degree \(\sum _i m_i\) is at most \(D\). Then there exist coefficients \(c : \mathcal{S}_{n,D} \to K\), indexed by the squarefree supports \(s \subseteq \mathrm{Fin}\, n\) with \(\left\lvert s\right\rvert \le D\), such that the squarefree polynomial \(\sum _{\left\lvert s\right\rvert \le D} c(s)\prod _{i \in s} X_i\) takes the same value as \(a\, X^m\) at every point \(x\) of the root cube \(\{ 1,\omega \} ^n\).
Let \(K\) be a finite field, let \(\omega \in K\) be nonzero, and write \(\{ 1,\omega \} ^n\) for the set of vectors in \(K^n\) each of whose coordinates equals \(1\) or \(\omega \). Suppose that every function \(f : \{ 1,\omega \} ^n \to K\) agrees, at every point of the cube, with some multilinear polynomial \(\sum _{S \subseteq \{ 1,\dots ,n\} } c_S \prod _{i \in S} X_i\); and suppose that some function on the cube cannot be approximated to within \(e\) errors by any polynomial of total degree at most \(\lfloor n/2 \rfloor + d\) — that is, for at least one \(f\), every such polynomial disagrees with \(f\) on more than \(e\) points of the cube. Then no polynomial \(P\) of total degree at most \(d\) agrees with the top monomial \(x \mapsto \prod _{i=1}^{n} x_i\) on all but at most \(e\) points of \(\{ 1,\omega \} ^n\).
Let \(K\) be a finite field and let \(\omega \in K\) satisfy \(\omega \neq 0\) and \(\omega \neq 1\), and fix natural numbers \(n, d, e, B\). Suppose that
and that
where \(\left\lvert \mathcal{S}_{n,\lfloor n/2\rfloor + d} \to K\right\rvert \) is the number of functions from \(\mathcal{S}_{n,\lfloor n/2\rfloor + d}\) to \(K\). Then there is no polynomial \(P \in K[X_1,\dots ,X_n]\) of total degree at most \(d\) whose bad count against the coordinate product \(x \mapsto \prod _{i} x_i\) on the root cube \(\{ 1,\omega \} ^{n}\) is at most \(e\); that is, every such \(P\) disagrees with \(\prod _i x_i\) at more than \(e\) points of \(\{ 1,\omega \} ^n\).
Let \(K\) be a finite field, let \(\omega \in K\) be nonzero, and fix natural numbers \(n, d, e, B\). Suppose that every function \(f : \{ 1,\omega \} ^n \to K\) agrees, at every point of the cube \(\{ 1,\omega \} ^n\), with some squarefree (multilinear) polynomial \(\sum _{S \subseteq \{ 1,\dots ,n\} } c_S \prod _{i \in S} X_i\). Suppose further that a finite family of polynomials \((P_c)_{c \in \mathrm{Cand}}\), indexed by a finite type \(\mathrm{Cand}\), is complete for total degree at most \(\lfloor n/2 \rfloor + d\), in the sense that every \(Q \in K[X_1,\dots ,X_n]\) with \(\deg Q \le \lfloor n/2 \rfloor + d\) agrees with some \(P_c\) at every point of the cube; that for each \(c\) the Hamming ball of radius \(e\) centered at the function \(x \mapsto P_c(x)\) contains at most \(B\) functions on the cube; and that the number of functions \(\{ 1,\omega \} ^n \to K\) strictly exceeds \(\left\lvert \mathrm{Cand}\right\rvert \cdot B\). Then there is no polynomial \(P\) with \(\deg P \le d\) such that \(P(x) \neq \prod _{i} x_i\) for at most \(e\) points \(x\) of the cube; that is, no polynomial of total degree at most \(d\) approximates the product \(\prod _i x_i\) on \(\{ 1,\omega \} ^n\) with at most \(e\) errors.
Let \(p\) be a prime and let \(a \in \mathbb {Z}/p\mathbb {Z}\). Then
That is, \(1 - a^{p-1}\) is the indicator of \(a = 0\) in the field \(\mathbb {Z}/p\mathbb {Z}\).
Let \(p\) and \(q\) be distinct primes and let \(\omega \in \mathbb {F}_{p^{q-1}}\) satisfy \(\omega ^q = 1\) and \(\omega \neq 1\). Suppose \(P\) is a polynomial in \(n + (q-1)\) variables over \(\mathbb {Z}/p\) of total degree at most \(d\) that disagrees with the \(\mathrm{MOD}_q\) function (the indicator, in \(\mathbb {Z}/p\), that the Hamming weight is divisible by \(q\)) on fewer than \(E\) Boolean inputs of length \(n + (q-1)\). Then there is a family \((R_r)_{r \in \{ 0,\dots ,q-1\} }\) of polynomials in \(n\) variables over \(\mathbb {F}_{p^{q-1}}\) such that each \(R_r\) has total degree at most \(d\) and disagrees with the residue-\(r\) indicator on at most \(E\) points of the cube \(\{ 1,\omega \} ^n\).
Let \(p\) and \(q\) be distinct primes, let \(\omega \in \mathbb {F}_{p^{q-1}}\) satisfy \(\omega ^q = 1\) and \(\omega \neq 1\), and let \(E\) and \(e\) be natural numbers with \(q \cdot E \le e\). If \(P\) is a polynomial in \(n + (q-1)\) variables over \(\mathbb {Z}/p\) of total degree at most \(d\) that disagrees with the \(\mathrm{MOD}_q\) function (the indicator that the Hamming weight is divisible by \(q\)) on fewer than \(E\) Boolean inputs, then there exists a polynomial \(Q\) in \(n\) variables over \(\mathbb {F}_{p^{q-1}}\) of total degree at most \(d\) that disagrees with the root product \(x \mapsto \prod _i x_i\) on at most \(e\) points of the cube \(\{ 1,\omega \} ^n\).
Let \(p\) and \(q\) be distinct primes and let \(\omega \in \mathbb {F}_{p^{q-1}}\) satisfy \(\omega ^q = 1\) and \(\omega \neq 1\). If \(P\) is a polynomial in \(n + (q-1)\) variables over \(\mathbb {Z}/p\) of total degree at most \(d\) that disagrees with the \(\mathrm{MOD}_q\) function (the indicator that the Hamming weight is divisible by \(q\)) on fewer than \(E\) Boolean inputs, then there exists a polynomial \(Q\) in \(n\) variables over \(\mathbb {F}_{p^{q-1}}\) of total degree at most \(d\) that disagrees with the root product \(x \mapsto \prod _i x_i\) on at most \(q \cdot E\) points of the cube \(\{ 1,\omega \} ^n\).
Let \(K\) be a field, \(\omega \in K\), let \(q\) be a positive natural number, let \(r \in \{ 0,\dots ,q-1\} \), and let \(x \in \{ 1,\omega \} ^n\). The associated padded Boolean input on \(n + (q-1)\) bits has as its first \(n\) bits the decoded bits of \(x\) (bit \(0\) where \(x_i = 1\), bit \(1\) elsewhere), while its final \(q-1\) bits consist of \((q - r) \bmod q\) ones followed by zeros.
Let \(p\) and \(q\) be primes, let \(\omega \in \mathbb {F}_{p^{q-1}}\) with \(\omega \neq 1\), and fix a residue \(r \in \{ 0,\dots ,q-1\} \). The map sending a root-cube point \(x \in \{ 1,\omega \} ^n\) to its padded Boolean input on \(n + (q-1)\) bits is injective.
Let \(p\) and \(q\) be primes, let \(\omega \in \mathbb {F}_{p^{q-1}}\) with \(\omega \neq 1\), let \(r \in \{ 0,\dots ,q-1\} \), and let \(x \in \{ 1,\omega \} ^n\). Write \(\mathrm{MOD}_q\) for the \(\mathbb {Z}/p\)-valued function on Boolean inputs that is \(1\) when the Hamming weight of its input is divisible by \(q\) and \(0\) otherwise. Then the image in \(\mathbb {F}_{p^{q-1}}\) of \(\mathrm{MOD}_q\) evaluated at the padded Boolean input associated to \((x, r)\) equals the residue-\(r\) indicator of \(x\): it is \(1\) exactly when the \(\omega \)-weight of \(x\) is congruent to \(r\) modulo \(q\), and \(0\) otherwise.
Let \(p\) and \(q\) be primes, let \(K = \mathbb {F}_{p^{q-1}}\), let \(\omega \in K\), let \(r \in \{ 0,\dots ,q-1\} \), and let \(P\) be a polynomial in \(n + (q-1)\) variables over \(\mathbb {Z}/p\). The residue-\(r\) restriction of \(P\) is the polynomial in \(n\) variables over \(K\) obtained by extending the coefficients of \(P\) along the canonical embedding \(\mathbb {Z}/p \to K\) and then substituting each variable according to the residue-\(r\) substitution: the first \(n\) variables become the affine root-cube decoders and the padding variables are fixed to the constants encoding residue \(r\).
Let \(p\) and \(q\) be primes, let \(\omega \in \mathbb {F}_{p^{q-1}}\) with \(\omega \neq 1\), let \(r \in \{ 0,\dots ,q-1\} \), let \(P\) be a polynomial in \(n + (q-1)\) variables over \(\mathbb {Z}/p\), and let \(x \in \{ 1,\omega \} ^n\). Evaluating the residue-\(r\) restriction of \(P\) at \(x\) gives the same value as evaluating \(P\) at the padded Boolean input associated to \((x, r)\) (with bits cast into \(\mathbb {Z}/p\)) and then mapping the result along the embedding \(\mathbb {Z}/p \to \mathbb {F}_{p^{q-1}}\).
Let \(p\) and \(q\) be primes, let \(\omega \in \mathbb {F}_{p^{q-1}}\), and let \(r \in \{ 0,\dots ,q-1\} \). For every polynomial \(P\) in \(n + (q-1)\) variables over \(\mathbb {Z}/p\), the residue-\(r\) restriction of \(P\) (a polynomial in \(n\) variables over \(\mathbb {F}_{p^{q-1}}\)) has total degree at most the total degree of \(P\).
Let \(p\) and \(q\) be primes, let \(K = \mathbb {F}_{p^{q-1}}\), let \(\omega \in K\), and let \(r \in \{ 0,\dots ,q-1\} \). The residue-\(r\) substitution assigns to each of the \(n + (q-1)\) variables a polynomial in \(n\) variables over \(K\): a variable \(X_j\) with \(j \lt n\) is sent to the affine decoder \((\omega - 1)^{-1}\, (X_j - 1)\), while the \(k\)-th padding variable (with \(k = j - n\)) is sent to the constant \(1\) if \(k \lt (q - r) \bmod q\) and to the constant \(0\) otherwise.
Let \(\iota \) be a finite index type with decidable equality, let \((\beta _j)_{j \in \iota }\) be a family of finite types, and fix a coordinate \(i \in \iota \). Let \(\mathrm{Bad}\) be a decidable predicate on \(\beta _i\), and let \(C\) be a natural number such that
Then the number of dependent functions \(f \in \prod _{j \in \iota } \beta _j\) whose value at \(i\) is bad satisfies
Let \(\iota \) be a finite index set, and for each \(i \in \iota \) let \(\beta _i\) be a finite set equipped with a predicate \(\mathrm{Bad}_i\) singling out its "bad" elements. Fix a natural number \(C\), and suppose that for every coordinate \(i\) the bad elements are sufficiently rare, in the sense that
Then, among all dependent tuples \(f\) in the product \(\prod _{i \in \iota } \beta _i\), those that are bad in at least one coordinate satisfy
For an index \(i\) with decidable equality on the index type, the equivalence
sending a dependent function to its value at \(i\) together with its restriction to the remaining coordinates.
Let \(\alpha \) be a finite type and \(\beta \) a finite nonempty type, let \(\mathrm{Bad} \subseteq \alpha \) be a finite set, let \(\mathrm{Fail} \subseteq \alpha \times \beta \) be a decidable relation, and let \(C\) be a natural number. Suppose that for every \(a \in \mathrm{Bad}\) we have \(\left\lvert \{ b \in \beta : \mathrm{Fail}(a,b)\} \right\rvert \cdot C \le \left\lvert \beta \right\rvert \). Then there exists a single \(b \in \beta \) such that \(\left\lvert \{ a \in \mathrm{Bad} : \mathrm{Fail}(a,b)\} \right\rvert \cdot C \le \left\lvert \mathrm{Bad}\right\rvert \).
Let \(\alpha \) and \(\beta \) be finite types, let \(P\) be a predicate on \(\alpha \) and \(Q\) a predicate on \(\alpha \times \beta \), and let \(C, B\) be natural numbers. Suppose that for every \(a \in \alpha \) satisfying \(P(a)\), the fiber \(\{ \, b \in \beta : Q(a,b)\, \} \) satisfies \(\left\lvert \{ \, b : Q(a,b)\, \} \right\rvert \cdot C \le B\). Then
Let \(\alpha \) and \(\beta \) be finite types, and let \(P\) be a decidable predicate on \(\alpha \). Then the number of pairs \((a,b) \in \alpha \times \beta \) whose first component satisfies \(P\) equals the number of elements \(a \in \alpha \) satisfying \(P\), multiplied by the cardinality \(\left\lvert \beta \right\rvert \).
Let \(q\) be a positive natural number. For a residue \(r \in \{ 0,\dots ,q-1\} \), the number of padding \(1\)-bits needed to shift Hamming-weight residue \(r\) to residue \(0\) modulo \(q\), namely \((q - r) \bmod q\): this is \(0\) when \(r = 0\) and \(q - r\) when \(r \gt 0\).
Let \(K\) be a finite field, let \(q\) be a prime, and let \(\omega \in K\) satisfy \(\omega ^q = 1\) and \(\omega \neq 1\). Given polynomials \(R_r\) in \(n\) variables over \(K\), one for each residue \(r \in \{ 0,\dots ,q-1\} \), each of total degree at most \(d\) and each disagreeing with the residue-\(r\) indicator on at most \(E\) points of \(\{ 1,\omega \} ^n\), there exists a polynomial \(Q\) in \(n\) variables over \(K\) of total degree at most \(d\) that disagrees with the root product \(x \mapsto \prod _i x_i\) on at most \(q \cdot E\) points of the cube.
Let \(K\) be a field and \(\omega \in K\). For a variable index \(i \in \{ 1,\dots ,n\} \), the affine polynomial \((\omega - 1)^{-1}\, (X_i - 1)\) in \(n\) variables over \(K\). When \(\omega \neq 1\) it decodes a root-cube coordinate as a Boolean value, sending \(1 \mapsto 0\) and \(\omega \mapsto 1\).
Let \(K\) be a field, let \(\omega \in K\) be a nonzero element, and let \(x \in \{ 1,\omega \} ^n\), so that each coordinate \(x_i\) equals either \(1\) or \(\omega \). Then for every coordinate \(i\),
Let \(K\) be a finite field and let \(\omega \in K\) with \(\omega \ne 1\). Then the root cube \(\{ 1,\omega \} ^n\), consisting of those vectors \(x : \mathrm{Fin}\, n \to K\) with \(x_i \in \{ 1,\omega \} \) for every coordinate \(i\), has exactly \(2^n\) elements.
Let \(K\) be a field and let \(\omega \in K\) with \(\omega \neq 0\). If \(x \in \{ 1,\omega \} ^n\), that is, \(x\) is a vector in \(K^n\) each of whose coordinates equals either \(1\) or \(\omega \), then every coordinate is nonzero: \(x_i \neq 0\) for each index \(i\).
Let \(K\) be a finite field, fix an element \(\omega \in K\), and let \(n, D, e, B\) be natural numbers. Write \(\{ 1,\omega \} ^n\) for the cube of vectors in \(\mathrm{Fin}\, n \to K\) whose every coordinate equals \(1\) or \(\omega \). Suppose a finite index set \(\mathrm{Cand}\) is equipped with a map assigning to each candidate \(c\) a polynomial \(P_c \in K[X_1,\dots ,X_n]\), and that the following hold: (i) every polynomial \(Q\) of total degree at most \(D\) agrees, at every point of \(\{ 1,\omega \} ^n\), with \(P_c\) for some candidate \(c\); (ii) for each candidate \(c\), the Hamming ball of radius \(e\) centered at the function \(x \mapsto P_c(x)\) contains at most \(B\) functions \(\{ 1,\omega \} ^n \to K\); and (iii) \(\left\lvert \{ 1,\omega \} ^n \to K\right\rvert \gt \left\lvert \mathrm{Cand}\right\rvert \cdot B\). Then it is not the case that every function \(f : \{ 1,\omega \} ^n \to K\) admits a polynomial \(Q\) of total degree at most \(D\) differing from \(f\) on at most \(e\) points of the cube.
Let \(K\) be a finite field, let \(\omega \in K\) satisfy \(\omega \neq 1\), and fix natural numbers \(n, D, e, B\). Suppose the ball bound
holds, together with the strict counting inequality
Then it is not the case that every function \(f : \{ 1,\omega \} ^n \to K\) admits a multivariate polynomial \(Q \in K[X_i : i \in \mathrm{Fin}\, n]\) of total degree at most \(D\) that disagrees with \(f\) on at most \(e\) points of the cube \(\{ 1,\omega \} ^n\); that is, some \(f\) fails to be approximated in this sense by any such \(Q\).
Let \(K\) be a finite field, let \(\omega \in K\) with \(\omega \ne 1\), and let \(n \in \mathbb {N}\). Writing \(\{ 1,\omega \} ^n\) for the set of vectors \(x : \mathrm{Fin}\, n \to K\) each of whose coordinates equals \(1\) or \(\omega \), the number of functions \(\{ 1,\omega \} ^n \to K\) is
Let \(K\) be a field, let \(\omega \in K\) be nonzero, and let \(x \in \{ 1,\omega \} ^n\), that is, a vector \(x \colon \mathrm{Fin}\, n \to K\) each of whose coordinates equals \(1\) or \(\omega \). Then for every subset \(s \subseteq \mathrm{Fin}\, n\),
where the first product runs over all coordinates and \(s^{c}\) denotes the complement of \(s\).
Let \(K\) be a finite field, let \(q\) be a prime, and let \(\omega \in K\) satisfy \(\omega ^q = 1\). For any family \((R_r)_{r \in \{ 0,\dots ,q-1\} }\) of polynomials in \(n\) variables over \(K\), the number of points of \(\{ 1,\omega \} ^n\) at which the combined polynomial \(\sum _{r} \omega ^{r} R_r\) differs from the root product \(\prod _i x_i\) is at most the sum over \(r\) of the number of points at which \(R_r\) differs from the residue-\(r\) indicator.
Let \(K\) be a finite field, let \(\omega \in K\), and let \(n, e \in \mathbb {N}\). Write \(\{ 1,\omega \} ^n\) for the root cube of vectors \(x : \{ 0,\dots ,n-1\} \to K\) each of whose coordinates equals \(1\) or \(\omega \), and consider functions \(f : \{ 1,\omega \} ^n \to K\), the Hamming distance between two such functions being the number of points of the cube on which they differ. For any function \(\mathrm{center} : \{ 1,\omega \} ^n \to K\), the number of functions whose Hamming distance from \(\mathrm{center}\) is at most \(e\) satisfies
where \(\left\lvert \{ 1,\omega \} ^n\right\rvert \) is the number of points of the cube and \(\left\lvert K\right\rvert \) is the number of elements of \(K\).
Let \(K\) be a field and \(\omega \in K\). The \(i\)-th decoded bit of a point \(x \in \{ 1,\omega \} ^n\) is the Boolean value that is \(0\) if \(x_i = 1\) and \(1\) otherwise. When \(\omega \neq 1\) this inverts the Boolean embedding, sending the coordinate value \(1\) to the bit \(0\) and the value \(\omega \) to the bit \(1\).
Let \(K\) be a finite field and fix \(\omega \in K\) with \(\omega \neq 0\), and let \(n\) be a natural number. Suppose that every function \(f\colon \{ 1,\omega \} ^n \to K\) agrees, at every point of the cube \(\{ 1,\omega \} ^n\), with some squarefree polynomial \(\sum _{S \subseteq \{ 1,\dots ,n\} } c_S \prod _{i \in S} X_i\) over \(K\). If a polynomial \(P\) in \(n\) variables over \(K\) has total degree at most \(d\) and disagrees with the top monomial \(x \mapsto \prod _{i} x_i\) on at most \(e\) points of \(\{ 1,\omega \} ^n\), then every function \(f\colon \{ 1,\omega \} ^n \to K\) admits a polynomial \(Q\) of total degree at most \(\lfloor n/2 \rfloor + d\) that disagrees with \(f\) on at most \(e\) points of the cube.
Let \(K\) be a field, let \(q\) be a prime, and let \(\omega \in K\) satisfy \(\omega ^q = 1\). For every point \(x \in \{ 1,\omega \} ^n\),
where \(w(x)\) denotes the number of coordinates of \(x\) equal to \(\omega \).
Let \(K\) be a field, \(\omega \in K\), and let \(q\) be a positive natural number. For a residue \(r \in \{ 0,\dots ,q-1\} \), the residue-\(r\) indicator is the \(K\)-valued function on the cube \(\{ 1,\omega \} ^n\) that takes the value \(1\) at \(x\) when the \(\omega \)-weight of \(x\) is congruent to \(r\) modulo \(q\), and \(0\) otherwise.
Let \(K\) be a field and \(\omega \in K\). The \(\omega \)-weight of a point \(x \in \{ 1,\omega \} ^n\) is the number of indices \(i\) with \(x_i = \omega \). Under the Boolean embedding \(0 \mapsto 1\), \(1 \mapsto \omega \), this is the Boolean Hamming weight.
Let \(F\) be a layered feedforward circuit over the Boolean alphabet \(\mathbb {F}_2 = \{ 0,1\} \) with input type \(\mathrm{Fin}\, n\), and suppose every node layer of \(F\) is finite. Then the size of \(F\) equals the sum of the cardinalities of its non-input layers:
that is, the total number of non-input nodes is obtained by adding up the numbers of nodes in layers \(1, \dots , \mathrm{depth}(F)\).
Let \(p\) and \(q\) be primes, and let \(F\) be a feedforward circuit over the alphabet \(\{ 0,1\} \) with input coordinates indexed by \(\mathrm{Fin}\, n\), finitely many nodes in each layer, and a single output node. Suppose every gate of \(F\) belongs to the \(\mathrm{AC}^0[p]\) gate set (identity, NOT, unbounded fan-in AND, and unbounded \(\mathrm{MOD}_p\) gates of every arity), and that \(F\) computes the \(\mathrm{MOD}_q\) function on every Boolean input \(x \in \{ 0,1\} ^n\). Let \(\ell , E \in \mathbb {N}\), and suppose the low-degree bad-count lower bound holds for \(\mathrm{MOD}_q\) (viewed as a function into \(\mathbb {Z}/p\)) with degree budget \(\bigl((p-1)\ell \bigr)^{\mathrm{depth}(F)}\) and count \(E\); that is, every polynomial in \((\mathbb {Z}/p)[X_1,\dots ,X_n]\) of total degree at most \(\bigl((p-1)\ell \bigr)^{\mathrm{depth}(F)}\) disagrees with \(\mathrm{MOD}_q\) on at least \(E\) points of the Boolean cube. Then
where \(\left\lvert F\right\rvert \) is the number of non-input nodes of \(F\).
Let \(p\) and \(q\) be primes, and let \(F\) be a layered feedforward circuit over the Boolean alphabet with input coordinates \(\mathrm{Fin}\, n\) and a single output, each of whose layers is finite. Suppose \(F\) uses only \(\mathrm{AC}^0[p]\) gates (the identity gate, the NOT gate, unbounded fan-in AND gates, and unbounded \(\mathrm{MOD}_p\) gates), and that \(F\) computes the \(\mathrm{MOD}_q\) function on its \(n\) Boolean inputs. Fix \(\ell , \delta \in \mathbb {N}\), and assume the relative low-degree bad-count bound: every polynomial \(P \in (\mathbb {Z}/p)[X_1,\dots ,X_n]\) of total degree at most \(\bigl((p-1)\ell \bigr)^{\mathrm{depth}(F)}\) disagrees with the \(\mathrm{MOD}_q\) target (viewed in \(\mathbb {Z}/p\)) on at least \(\delta \cdot 2^{n}\) of the \(2^{n}\) Boolean inputs. Then the size of \(F\) satisfies \(\delta \cdot 2^{\ell } \le \left\lvert F\right\rvert \).
Let \(p\) and \(q\) be distinct primes, write \(K = \mathbb {F}_{p^{q-1}}\), and let \(d, E, e, B\) be natural numbers with \(q \cdot E \le e\). Assume the counting bounds
where \(N\) is the number of \(K\)-valued functions on the collection of subsets \(S \subseteq \{ 1,\dots ,n\} \) with \(\left\lvert S\right\rvert \le \lfloor n/2 \rfloor + d\). Then every polynomial in \(n + (q-1)\) variables over \(\mathbb {Z}/p\) of total degree at most \(d\) disagrees with the \(\mathrm{MOD}_q\) function (the indicator that the Hamming weight is divisible by \(q\)) on at least \(E\) Boolean inputs.
Let \(K\) be a field, let \(\omega \in K\) be nonzero, and let \((c_S)_{S \subseteq \mathrm{Fin}\, n}\) be a family of coefficients in \(K\) indexed by the subsets of \(\mathrm{Fin}\, n\). Then there exist polynomials \(P_1, P_2 \in K[X_i : i \in \mathrm{Fin}\, n]\), each of total degree at most \(\lfloor n/2 \rfloor \), such that for every point \(x \in \{ 1,\omega \} ^n\) (that is, every \(x\) with \(x_i \in \{ 1,\omega \} \) for all \(i\)),
Let \(K\) be a field, fix a scalar \(\omega \in K\) with \(\omega \neq 0\), and let \(n\) be a natural number. For any family of coefficients \((c_S)\) indexed by the subsets \(S \subseteq \{ 1,\dots ,n\} \), form the multilinear polynomial \(F = \sum _{S} c_S \prod _{i \in S} X_i\) in \(K[X_1,\dots ,X_n]\). Then there exist polynomials \(P_1, R \in K[X_1,\dots ,X_n]\), each of total degree at most \(\lfloor n/2 \rfloor \), such that for every point \(x \in \{ 1,\omega \} ^n\) (that is, every \(x\) with each coordinate \(x_i\) equal to \(1\) or to \(\omega \)) one has
Let \(K\) be a field, let \(n\) be a natural number, and let \(s \subseteq \mathrm{Fin}\, n\). For every point \(x : \mathrm{Fin}\, n \to K\), the squarefree monomial \(\prod _{i \in s} X_i\) evaluated at \(x\) equals \(\prod _{i \in s} x_i\).
Fix a field \(K\) and let \(s \subseteq \{ 0,1,\dots ,n-1\} \) be a set of variable indices. The total degree of the squarefree monomial \(\prod _{i \in s} X_i\) in the polynomial ring \(K[X_0,\dots ,X_{n-1}]\) is at most \(\left\lvert s\right\rvert \).
Given a circuit \(F\) using only \(\mathrm{AC}^0[p]\) gates and a LayerPolyFamily at layer \(d\), this produces a LayerPolyFamily at layer \(d+1\). Its seed type is the product of the previous seed type with one gate seed per node of layer \(d+1\), and each node polynomial is the gate approximator applied to the layer-\(d\) polynomials of its inputs.
Let \(p\) be a prime and let \(v = (v_i)_{i \in \{ 1,\dots ,n\} }\) be a nonzero tuple of elements of the field \(\mathbb {Z}/p\mathbb {Z}\). Then at most half of the \(2^n\) subsets of the index set \(\{ 1,\dots ,n\} \) have vanishing subset sum; that is, twice the number of subsets \(s \subseteq \{ 1,\dots ,n\} \) for which \(\sum _{i \in s} v_i = 0\) is at most \(2^n\).
Let \(\iota \) be a finite index type, let \(\beta \) be a finite type, and let \(P\) be a decidable predicate on \(\beta \). Then the number of functions \(f : \iota \to \beta \) such that \(P(f(i))\) holds for every \(i \in \iota \) equals \(\left\lvert \{ \, b \in \beta : P(b)\, \} \right\rvert ^{\left\lvert \iota \right\rvert }\).
Let \(D\) be a tournament on a finite, nonempty vertex set \(V\) — a directed graph with no loops in which, for every pair of distinct vertices \(u\) and \(v\), exactly one of the arcs \(u \to v\) and \(v \to u\) is present. Then \(D\) has a directed Hamilton path: there exist vertices \(u, v\) and a directed path from \(u\) to \(v\) that visits every vertex of \(V\) exactly once.
Let \(N \ge 1\), let \(\eta \gt 0\), let \(\ell \) be a loss sequence for \(N\) experts over \(T\) rounds, and let \(c \in \mathbb {R}\). Suppose that for every round \(t \lt T\) the log potential satisfies the one-step bound
where \(\mathrm{hedgeLoss}_t\) is the expected loss of Hedge in round \(t\). Then the regret of Hedge is bounded by
Consider a finite two-player zero-sum game with \(M \ge 1\) row actions and \(N\) column actions and payoff matrix \(A\), every entry of which satisfies \(0 \le A_{ij} \le 1\), played over \(T \ge 1\) rounds. Suppose that in each round \(t\) the row player commits to a probability distribution \(p_t\) over the row actions (so \(\sum _i p_t(i) = 1\)) while the column player chooses an action \(j_t\), and that the cumulative loss regret is bounded by \(R\):
where the minimum runs over the row actions. Then the row player’s cumulative expected payoff satisfies
with the maximum taken over the row actions.
Fix \(n\) and a parameter \(p\) with \(0 \le p \le 1\), let \(U \subseteq \mathrm{Fin}\, n\) be a set of coordinates, and let \(k \ge 1\) be an integer satisfying \(3k \le p\, \left\lvert U\right\rvert \). Draw a Bernoulli(\(p\)) random restriction \(\rho \) on \(n\) variables, and let \(J\) denote its set of free variables. Then
Let \(p \in \mathbb {R}\) satisfy \(0 \le p \le 1\), let \(U \subseteq \mathrm{Fin}\, n\) be a finite set of coordinates, and let \(k\) be a positive integer with \(3k \le p\, \left\lvert U\right\rvert \). Writing \(J\) for the set of free variables of a Bernoulli(\(p\))-random restriction \(\rho \) on \(n\) variables, the probability that fewer than \(k\) coordinates of \(U\) remain free satisfies
Fix \(n\), a real parameter \(p\), and two subsets \(U, S \subseteq [n]\). Let \(\rho \) be a Bernoulli(\(p\)) random restriction on \(n\) variables, meaning each restriction is drawn with weight \(p^{\left\lvert J\right\rvert }\bigl(\tfrac {1-p}{2}\bigr)^{\, n-\left\lvert J\right\rvert }\), where \(J\) denotes its set of free variables. Then the probability that the free variables lying inside \(U\) are exactly those of \(S\) is
Let \(p \in \mathbb {R}\) with \(0 \le p \le 1\), and let \(E\) be a decidable event on restrictions of \(n\) variables, where a restriction assigns each coordinate either a fixed Boolean value or leaves it free. Under the Bernoulli(\(p\))-random restriction — each restriction \(\rho \) with free-variable set \(J\) carries weight \(p^{\left\lvert J\right\rvert }\bigl(\tfrac {1-p}{2}\bigr)^{\, n-\left\lvert J\right\rvert }\) — the probability of \(E\), namely the weighted sum of the indicator of \(E\) over all restrictions, is nonnegative: \(0 \le \Pr _\rho [E]\).
Fix \(n\) and a parameter \(p\) with \(0 \le p \le 1\), and consider the Bernoulli(\(p\)) random restriction on \(n\) variables, under which each restriction \(\rho \) receives weight \(p^{k}\, \bigl(\tfrac {1-p}{2}\bigr)^{\, n-k}\), where \(k\) is the number of free variables of \(\rho \). For any decidable predicate \(E\) on restrictions, writing \(\Pr [E]\) for the total weight of the restrictions satisfying \(E\), the probability that \(E\) fails equals one minus the probability that \(E\) holds:
Fix \(n \in \mathbb {N}\), a real parameter \(p\), and a subset \(T \subseteq \{ 1, \dots , n\} \) of the coordinates. Draw a Bernoulli(\(p\)) random restriction \(\rho \) on \(n\) variables, and write \(J\) for its set of free variables (the coordinates \(\rho \) leaves unset). Then the probability that every coordinate of \(T\) is left free is
that is, the total Bernoulli weight of those restrictions whose free variables include \(T\) equals \(p^{\left\lvert T\right\rvert }\).
Let \(\rho \) be a restriction on \(n\) variables, and let \(U \subseteq \{ 1,\dots ,n\} \) be a set of indices. Then the number of indices in \(U\) at which \(\rho \) is free (unset), viewed as a real number, equals
where \(J\) is the set of free variables of \(\rho \) and each summand is \(1\) when \(\rho (i)\) is unset and \(0\) otherwise.
Let \(\rho \) be a restriction on \(n\) variables, let \(J\subseteq [n]\) be its set of free coordinates, and let \(U\subseteq [n]\) and \(x\in \{ 0,1\} ^n\) be arbitrary. Write \(\hat x\) for the extension of \(x\) by \(\rho \), namely the point of \(\{ 0,1\} ^n\) that agrees with \(x\) on the free coordinates and takes the fixed value of \(\rho \) on the remaining ones. Then the Walsh character \(\chi _U\) evaluated at \(\hat x\) factors as
where \(b_i\) denotes the Boolean value that \(\rho \) fixes at coordinate \(i\) and \(\sigma \) is the sign encoding sending \(\mathtt{false}\mapsto 1\) and \(\mathtt{true}\mapsto -1\). Thus the character splits into a Walsh character in the free coordinates lying in \(U\) times a constant \(\pm 1\) sign contributed by the fixed coordinates of \(U\).
Let \(C\) be a Boolean circuit on \(n \ge 1\) input bits with depth at most \(d\), size at most \(s\), and gate fan-in at most \(w\), where \(d \ge 2\) and \(s, w \ge 1\). Let \(l \ge 1\) and \(t\) be natural numbers, and let \(\varepsilon \gt 0\) be a real number such that the switching parameters satisfy \(s \cdot 2^{-l} \le \varepsilon /2\) and \(2^{-t} \le \varepsilon /2\). Write \(\delta = \frac{1}{40w}\bigl(\frac{1}{40l}\bigr)^{d-2}\) for the composed restriction rate, \(\sigma \) for the sign encoding of a bit (\(\sigma (\mathtt{false}) = 1\), \(\sigma (\mathtt{true}) = -1\)), and \(\widehat{\sigma \circ C}(U)\) for the Fourier–Walsh coefficient at \(U \subseteq [n]\) of the \(\pm 1\)-encoding of the function computed by \(C\). Then the Fourier weight above degree \(3(t+1)/\delta \) is small:
Let \(C\) be a Boolean circuit on \(n \ge 1\) input bits with depth at most \(d\), size at most \(s\), and gate fan-in at most \(w\), where \(d \ge 2\) and \(s, w \ge 1\), and let \(\varepsilon \) be a real number with \(0 \lt \varepsilon \le 1\). Set \(l = \lceil \log _2(2s/\varepsilon ) \rceil \) and \(t = \lceil \log _2(2/\varepsilon ) \rceil \), and write \(\delta = \frac{1}{40w}\bigl(\frac{1}{40l}\bigr)^{d-2}\), \(\sigma \) for the sign encoding of a bit (\(\sigma (\mathtt{false}) = 1\), \(\sigma (\mathtt{true}) = -1\)), and \(\widehat{\sigma \circ C}(U)\) for the Fourier–Walsh coefficient at \(U \subseteq [n]\) of the \(\pm 1\)-encoding of the function computed by \(C\). Then
that is, the Fourier spectrum of \(\sigma \circ C\) is concentrated — up to Chernoff tail terms that vanish as \(n \to \infty \) — below degree \(3\bigl(\lceil \log _2(2/\varepsilon )\rceil + 1\bigr) \cdot 40 w \, \bigl(40\lceil \log _2(2s/\varepsilon )\rceil \bigr)^{d-2} = O\! \left(w \log (s/\varepsilon )^{d-2} \log (1/\varepsilon )\right)\).
Let \(w \ge 1\), \(l \ge 1\), and \(d\) be natural numbers, let \(a\) be a real number, and let \(m\) be a natural number. Writing \(\delta = \frac{1}{40w}\bigl(\frac{1}{40l}\bigr)^{d-2}\), where the exponent \(d-2\) is taken to be \(0\) when \(d \le 2\), the inequality \(a \le \delta m\) holds if and only if
(Taking \(a = 3(t+1)\) with \(t \approx \log _2(1/\varepsilon )\) and \(l \approx \log _2(s/\varepsilon )\), this exhibits the LMN degree threshold as \(O\! \left(w \log (s/\varepsilon )^{d-2} \log (1/\varepsilon )\right)\).)
Fix \(n\) and a real number \(p\), and let \(\rho \) range over all restrictions on the \(n\) variables, each weighted by its Bernoulli weight \(w_p(\rho )\). Writing \(J(\rho )\) for the set of free variables of \(\rho \), then for every set \(U\) of variables,
Equivalently, when \(0 \le p \le 1\) so that \(w_p\) is a probability distribution, the expected number of free variables landing in \(U\) is \(p\, \left\lvert U\right\rvert \).
Fix \(n \in \mathbb {N}\), a parameter \(p \in \mathbb {R}\), and a set of coordinates \(U \subseteq \mathrm{Fin}\, n\). For a restriction \(\rho \) on \(n\) variables write \(J(\rho )\) for its set of free variables. Summing over all such restrictions, each weighted by its Bernoulli weight, one has
Let \(f:\{ 0,1\} ^n\to \mathbb {R}\) be a Boolean function, let \(S\subseteq [n]\) be a frequency set, and let \(p\in \mathbb {R}\). For a restriction \(\rho \) on \(n\) variables, write \(f_\rho \) for the restricted function and \(k(\rho )\) for the number of free variables of \(\rho \). Summing the Fourier coefficient \(\widehat{f_\rho }(S)\) over all restrictions \(\rho \), each weighted by its Bernoulli weight \(p^{k(\rho )}\bigl(\tfrac {1-p}{2}\bigr)^{\, n-k(\rho )}\), gives
(When \(0\le p\le 1\) these weights form the distribution of a Bernoulli\((p)\)-random restriction, and the left-hand side is the expectation \(\mathbb {E}_\rho [\widehat{f_\rho }(S)]\).)
Fix a dimension \(n\), a real number \(p\), a Boolean function \(f:\{ 0,1\} ^n\to \mathbb {R}\), and a set \(S\subseteq [n]\). For a restriction \(\rho \) on the \(n\) variables, let \(k(\rho )\) denote its number of free variables and give it the Bernoulli weight \(w_p(\rho )=p^{\, k(\rho )}\bigl(\tfrac {1-p}{2}\bigr)^{\, n-k(\rho )}\); write \(f_\rho \) for the restriction of \(f\) by \(\rho \) and \(\hat g(T)\) for the Fourier–Walsh coefficient of a Boolean function \(g\) at a set \(T\). Then, with the left sum over all restrictions on the \(n\) variables and the right sum over all \(U\subseteq [n]\),
When \(0\le p\le 1\) the weights \(w_p(\rho )\) form a probability distribution, so the left-hand side is the expectation of \(\widehat{f_\rho }(S)^2\) over a Bernoulli\((p)\) random restriction.
Fix a real parameter \(p\) and a coordinate \(i\) among the \(n\) variables. Summing the Bernoulli weight of \(\rho \) over all restrictions \(\rho \) on \(n\) variables that leave \(i\) free — that is, those for which \(i\) is a free variable of \(\rho \) — gives
where \(F(\rho )\) denotes the set of free variables of \(\rho \).
Fix \(n\) and two distinct variables \(i \neq j\) among the \(n\) variables, and let \(p \in \mathbb {R}\). Write \(w_p(\rho )\) for the Bernoulli weight of a restriction \(\rho \) on \(n\) variables, namely \(p^{k}\left(\tfrac {1-p}{2}\right)^{n-k}\) where \(k\) is the number of free variables of \(\rho \). Then, summing over all such restrictions, the total weight carried by those \(\rho \) that leave both \(i\) and \(j\) free equals \(p^2\); that is,
Let \(f : \{ 0,1\} ^n \to \mathbb {R}\), and let \(\rho \) be a restriction on \(n\) variables; write \(J \subseteq [n]\) for its set of free coordinates, namely the coordinates left unset by \(\rho \). Let \(f_\rho : \{ 0,1\} ^n \to \mathbb {R}\) be the restricted function whose value at a point \(x\) is obtained by evaluating \(f\) at the assignment that agrees with \(x\) on the free coordinates in \(J\) and takes the bit fixed by \(\rho \) on every other coordinate. Then for every \(S \subseteq [n]\),
where \(\sigma \) denotes the sign encoding and, for each \(i \in U \setminus J\) (which lies outside \(J\) and is therefore a coordinate fixed by \(\rho \)), \(\rho _i \in \{ \mathtt{false},\mathtt{true}\} \) is the bit that \(\rho \) assigns there. This is O’Donnell’s Corollary 3.22.
Let \(U, S \subseteq [n]\) with \(S \subseteq U\), and let \(\rho \) be a restriction on \(n\) variables whose free-variable set is \(J\). Write \(\sigma (b) \in \{ \pm 1\} \) for the sign encoding of a bit \(b\), and let \(\ell (U,S,i,\rho (i))\) denote the per-coordinate local factor at coordinate \(i\). Then
where the left-hand side is read as \(0\) whenever \(U \cap J \neq S\). Equivalently, the product on the right vanishes unless the free coordinates of \(\rho \) lying in \(U\) are exactly \(S\), in which case it equals the product of the \(\pm 1\)-encodings of the bits that \(\rho \) fixes on \(U \setminus J\).
Let \(\rho \) be a restriction on \(n\) variables — a map assigning each index \(i \in \{ 0,\dots ,n-1\} \) either a fixed Boolean value or the free mark — and let \(T \subseteq \{ 0,\dots ,n-1\} \). Then the indicator of the event that every index in \(T\) is free under \(\rho \) factors as a product over all \(n\) coordinates:
where the factor \(c_i\) equals \(\mathbf{1}[\rho (i)\ \text{is free}]\) when \(i \in T\) and equals \(1\) otherwise.
Let \(g : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean-valued function, and let \(\sigma \) denote the sign encoding of a bit, \(\sigma (\mathtt{false}) = 1\) and \(\sigma (\mathtt{true}) = -1\). Then the real-valued function \(x \mapsto \sigma (g(x))\) takes only the values \(1\) and \(-1\) at every point \(x\) of the cube.
Given sets \(U, S \subseteq [n]\), a coordinate \(i\) and a per-coordinate value \(v\), the local factor is \(1\) if \(i \in S\) and \(v = \mathtt{none}\) (the coordinate must be free), \(0\) if \(i \in S\) and \(v\) is fixed; for \(i \in U \setminus S\) it is \(0\) when \(v = \mathtt{none}\) and \(\mathrm{boolToSign}(b)\) when \(v = \mathtt{some}\ b\); and \(1\) for all remaining coordinates. It encodes the event \(U \cap \mathrm{freeVars} = S\) together with the sign contributed on \(U \setminus S\).
Let \(\rho : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\) be a restriction on \(n\) variables, and let \(i \in \mathrm{Fin}\, n\) be a coordinate. Then \(i\) belongs to the set of free variables of \(\rho \) if and only if \(\rho (i)\) is unset; that is, exactly when \(\rho \) leaves the coordinate \(i\) free rather than fixing it to a Boolean value.
This is O’Donnell’s Lemma 4.21, with constant \(4\) in place of \(3\). Let \(f : \{ 0,1\} ^n \to \{ 0,1\} \) be a Boolean function, let \(p \in \mathbb {R}\) with \(0 \le p \le 1\), and let \(k \in \mathbb {N}\) with \(k \ge 1\). Let \(\rho \) be a Bernoulli(\(p\))-random restriction on \(n\) variables: each coordinate is independently left free with probability \(p\) and otherwise fixed to a uniformly random bit, so a restriction with free-variable set \(J\) has probability \(p^{\left\lvert J\right\rvert }\bigl(\tfrac {1-p}{2}\bigr)^{\, n-\left\lvert J\right\rvert }\). Write \(f_\rho \) for the function obtained from \(f\) by fixing the coordinates outside \(J\) according to \(\rho \), \(\mathrm{DT}(g)\) for the least depth of a decision tree computing a Boolean function \(g\), and \(\sigma \) for the sign encoding (\(\sigma (\mathtt{false}) = 1\), \(\sigma (\mathtt{true}) = -1\)). Setting \(\varepsilon = \Pr _\rho \bigl[\mathrm{DT}(f_\rho ) \ge k\bigr]\), the Fourier spectrum of the \(\pm 1\)-encoding \(\sigma \circ f\) is \(4\varepsilon \)-concentrated below degree \(3k/p\):
where the degree condition is stated division-free as \(3k \le p\, \left\lvert U\right\rvert \).
Let \(n\) be a natural number, let \(p \in \mathbb {R}\), and let \(S, U \subseteq \{ 1, \dots , n\} \). For each coordinate \(i\), assign the weight \(p\) if \(i \in S\), the weight \(1 - p\) if \(i \in U \setminus S\), and the weight \(1\) otherwise. Then the product of these weights over all \(n\) coordinates equals
Given \(f : \{ 0,1\} ^n \to \mathbb {R}\) and a restriction \(\rho \), the restricted function \(f_\rho = \mathrm{restrictBF}\, f\, \rho \) is defined by
so free coordinates are read from the input \(x\) and fixed coordinates from \(\rho \). It is the real-valued analogue of \(\mathrm{restrictFn}\).
Write \(\sigma \colon \{ 0,1\} \to \{ -1,1\} \) for the sign encoding, \(\sigma (0)=1\) and \(\sigma (1)=-1\), and for a function \(g\) on the cube (Boolean- or real-valued) let \(g_\rho \) denote its restriction along \(\rho \), obtained by reading each free coordinate from the input and each fixed coordinate from \(\rho \). Let \(f\colon \{ 0,1\} ^n\to \{ 0,1\} \) be a Boolean-valued function and let \(\rho \) be a restriction on \(n\) variables. Then restricting the real-valued function \(x\mapsto \sigma (f(x))\) along \(\rho \) agrees with first restricting \(f\) and then applying the sign encoding; that is, \((\sigma \circ f)_\rho =\sigma \circ f_\rho \), so that for every point \(x\),
For a restriction \(\rho \) and a set \(T \subseteq [n]\),
the product of the \(\pm 1\)-encodings of the bits that \(\rho \) fixes on \(T\). On coordinates where \(\rho \) is free the default value \(\mathtt{false}\) is used; this junk value is harmless because the definition is only ever paired with indicators forcing \(T\) to consist of fixed coordinates.
Let \(\rho \) be a restriction on \(n\) variables and let \(T \subseteq [n]\), and write \(s(\rho , T) \in \mathbb {R}\) for the sign product of the bits that \(\rho \) fixes on \(T\), that is, the product over \(i \in T\) of the \(\pm 1\)-encoding of \(\rho \)’s value at \(i\). Being a product of values in \(\{ -1, 1\} \), it satisfies
Fix \(n \in \mathbb {N}\) and a real parameter \(p\), and recall that a restriction on \(n\) variables is a map \(\rho : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\) assigning each variable either a fixed Boolean value or the free marker \(\mathrm{none}\). Write \(w_p(\rho ) = p^{\, k}\bigl(\tfrac {1-p}{2}\bigr)^{\, n-k}\) for its Bernoulli weight, where \(k\) is the number of free variables of \(\rho \), and let \(\omega _p\) be the per-variable weight given by \(\omega _p(\mathrm{none}) = p\) and \(\omega _p(v) = \tfrac {1-p}{2}\) for a fixed value \(v\). Then for every family \(h_i : \mathrm{Option}\, \mathrm{Bool} \to \mathbb {R}\) indexed by \(i \in \mathrm{Fin}\, n\),
where the outer sum ranges over all restrictions \(\rho \) on \(n\) variables and each inner sum ranges over \(v \in \mathrm{Option}\, \mathrm{Bool}\).
Fix \(n\), a real parameter \(p\), a subset \(U \subseteq [n]\), and a threshold \(k \in \mathbb {N}\). For a restriction \(\rho \) on \(n\) variables write \(J\) for its set of free variables, and write \(\Pr _\rho [\cdot ]\) for the sum, over all restrictions, of the Bernoulli weight \(p^{\left\lvert J\right\rvert }\bigl(\tfrac {1-p}{2}\bigr)^{\, n-\left\lvert J\right\rvert }\) times the indicator of the event (a genuine probability when \(0 \le p \le 1\)). Then
The identity holds for every real \(p\); no bound on \(p\) is assumed.
Fix \(n \in \mathbb {N}\), and for each coordinate \(i \in \{ 1,\dots ,n\} \) let \(h_i : \{ \mathtt{true},\mathtt{false},\ast \} \to \mathbb {R}\) be a real-valued weight assigned to the three possible states of that coordinate (fixed to true, fixed to false, or left free). Then summing the product of these weights over all \(3^n\) restrictions \(\rho \) on \(n\) variables equals the product over coordinates of the coordinatewise total weights:
where \(\rho \) ranges over all restrictions on \(n\) variables and \(v\) ranges over the three per-coordinate states.
Fix \(n\), a real parameter \(p\), a function \(F : \{ 0,1\} ^n \to \mathbb {R}\), and \(k \in \mathbb {N}\). For a restriction \(\rho \) on \(n\) variables with free-variable set \(J\), write \(w_p(\rho ) = p^{\left\lvert J\right\rvert }\bigl(\tfrac {1-p}{2}\bigr)^{\, n-\left\lvert J\right\rvert }\) for its Bernoulli weight, \(F_\rho \) for the restriction of \(F\) by \(\rho \) (free coordinates read from the input, fixed coordinates from \(\rho \)), \(\hat g(T)\) for the Fourier–Walsh coefficient of a function \(g\) at a set \(T\), and \(\Pr _\rho [\cdot ]\) for the \(w_p\)-weighted sum of an event’s indicator over all restrictions. Then
When \(0 \le p \le 1\) the left-hand side is \(\sum _{\left\lvert S\right\rvert \ge k} \mathbb {E}_\rho \bigl[\widehat{F_\rho }(S)^2\bigr]\), so this is O’Donnell’s Proposition 4.17 summed over the high-degree frequency sets.
Fix a real parameter \(p\), two coordinate sets \(U, S \subseteq [n]\), and a coordinate \(i \in [n]\). Writing \(w_p\) for the per-variable Bernoulli weight — which assigns \(p\) to the free value and \((1-p)/2\) to each of the two fixed values — and \(\ell _{U,S,i}\) for the per-coordinate local factor at \(i\), the weighted sum over the three possible per-coordinate values \(v\) (free, or fixed to either Boolean) is
Fix a real parameter \(p\), subsets \(U, V, S \subseteq [n]\), and a coordinate \(i \in [n]\). Averaging over the three per-coordinate values \(v\) — the free value and the two fixed signs — with the per-variable weight \(w_p(v)\), the product of the two local factors, one for the pair \((U,S)\) and one for the pair \((V,S)\), satisfies
where \(w_p\) denotes the per-variable weight and \(\ell _{W,S}\) the local factor determined by the sets \(W\) and \(S\).
Let \(0 \le p \le 1\), and let \(U \subseteq \mathrm{Fin}\, n\) be a set of coordinates. For a restriction \(\rho \) on \(n\) variables, write \(w_p(\rho )\) for its Bernoulli weight and \(F(\rho )\) for its set of free variables. Then the \(w_p\)-weighted average, over all restrictions \(\rho \) on \(n\) variables, of the squared deviation of \(\left\lvert U \cap F(\rho )\right\rvert \) from \(p\left\lvert U\right\rvert \) satisfies
Fix a prime \(p\), a natural number \(n\), and a subset \(C \subseteq \{ 0,1,\dots ,n-1\} \) of coordinates, and let \(V_C\) denote the support submodule of \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) consisting of vectors whose support lies in \(C\). Let \(E \colon (C \to \mathbb {F}_p) \times (C \to \mathbb {F}_p) \to V_C\) be the extension map, which sends a pair \((f,g)\) to the vector agreeing with \(f\) and \(g\) on \(C\) and equal to zero off \(C\), and let \(R \colon V_C \to (C \to \mathbb {F}_p) \times (C \to \mathbb {F}_p)\) be the restriction map, which sends a vector supported on \(C\) to its two component functions restricted to \(C\). Then \(R \circ E\) is the identity: for every pair \(x = (f,g)\) of functions \(f, g \colon C \to \mathbb {F}_p\) one has \(R(E(x)) = x\).
Let \(\alpha \) be a linearly ordered type. For every \(x \in \alpha \) and every finite list \(\mathit{xs}\) over \(\alpha \), the value returned by the timed insertion of \(x\) into \(\mathit{xs}\) (discarding the accumulated time cost) coincides with the ordinary ordered insertion of \(x\) into \(\mathit{xs}\) with respect to the order \(\le \): the list obtained by placing \(x\) immediately before the first entry \(y\) of \(\mathit{xs}\) satisfying \(x \le y\), and at the end if no such entry exists.
Let \(\alpha \) be a linearly ordered type and let \(xs\) be a list of elements of \(\alpha \). The value returned by the timed insertion sort of \(xs\) equals the ordinary insertion sort of \(xs\) taken with respect to the order relation \(\le \).
Let \(G\) be a finite simple graph on a nonempty vertex set, and suppose its edge connectivity satisfies \(\kappa '(G) \ge 2\). Then \(G\) has a diconnected orientation: one can assign a direction to each edge of \(G\), forming a digraph \(D\) whose underlying graph is \(G\), so that for every ordered pair of vertices \(u, v\) there is a directed path in \(D\) from \(u\) to \(v\).
Let \((\Omega ,\mu )\) be a probability space, and let \(A\) be a random \(k\times d\) real matrix, that is, a measurable map assigning to each \(\omega \in \Omega \) a matrix \(A(\omega )\in \mathbb {R}^{k\times d}\). Suppose the rows of \(A\) are mutually independent, meaning the \(k\) random vectors \(\omega \mapsto \bigl(A(\omega )_{i,1},\dots ,A(\omega )_{i,d}\bigr)\in \mathbb {R}^{d}\), for \(i=1,\dots ,k\), are mutually independent under \(\mu \). Then for every fixed \(x\in \mathbb {R}^{d}\) the \(k\) coordinates of the product \(A(\omega )\, x\), namely the real-valued random variables \(\omega \mapsto (A(\omega )\, x)_i=\sum _{j=1}^{d}A(\omega )_{i,j}\, x_j\) for \(i=1,\dots ,k\), are mutually independent under \(\mu \).
(Roy, 1967; Gallai, 1968.) Let \(D\) be a digraph on a finite, nonempty vertex set \(V\), and suppose the underlying simple graph of \(D\) has chromatic number \(k\). Then \(D\) contains a directed path — a directed walk \(v_0, v_1, \dots , v_\ell \) whose consecutive vertices are joined by arcs \(v_{i-1} \to v_i\) and whose vertices are pairwise distinct — of length at least \(k - 1\), that is, with \(\ell \ge k - 1\) arcs.
Let \(p\) be a prime and let \(c \in \mathbb {N}\) be arbitrary. Then for every \(x \in \mathbb {Z}/p\mathbb {Z}\),
This is the form of Fermat’s little theorem underlying RSA, and it holds for all \(x\), including \(x = 0\).
Let \((n,e)\) be an RSA public key belonging to a secret key that also supplies distinct primes \(p\) and \(q\) with \(n = pq\), a private exponent \(d\), and a natural number \(k\) satisfying \(ed = 1 + k(p-1)(q-1)\). Then for every message \(m \in \mathbb {N}\), decrypting the encryption of \(m\) returns the class of \(m\) modulo \(n\): in \(\mathbb {Z}/n\mathbb {Z}\) we have \(\bigl(\bar m^{\, e}\bigr)^{d} = \bar m\).
Let \(p, q, m, ed \in \mathbb {N}\) with \(p\) and \(q\) coprime, and write \(\bar m\) for the residue class of \(m\) in each relevant quotient ring. If \(\bar m^{\, ed} = \bar m\) in \(\mathbb {Z}/p\mathbb {Z}\) and \(\bar m^{\, ed} = \bar m\) in \(\mathbb {Z}/q\mathbb {Z}\), then \(\bar m^{\, ed} = \bar m\) in \(\mathbb {Z}/pq\mathbb {Z}\).
Let \(p\) be a prime, and let \(m\), \(N\), and \(c\) be natural numbers satisfying \(N = 1 + c(p-1)\). Then
that is, \(m^{N}\) and \(m\) have the same image in \(\mathbb {Z}/p\mathbb {Z}\). (The congruence holds for every \(m\), including multiples of \(p\).)
The inductive type \(\mathtt{AuxVar}\, V\) extends a variable type \(V\) with two constructors: \(\mathtt{orig}\, v\) wraps an original variable \(v : V\), and \(\mathtt{extra}\, i\, j\) introduces a fresh auxiliary variable \(y_{i,j}\) used as a chain link when encoding clause \(i\) of length greater than three.
Given a clause index \(c\_ idx\), a lifting map \(ml\), a suffix list of literals \(\mathit{lits}\), and a chain counter \(j \ge 1\), SATTo3SAT.buildChain produces the middle and final 3-clauses of the chain encoding. Each middle step emits \(\langle \lnot y_{j-1},\, l_i,\, y_j\rangle \); the final step emits \(\langle \lnot y_{j-1},\, l_{n-1},\, l_n\rangle \) when only two literals remain.
Fix a clause index \(i\) and a clause presented as a list of literals over a variable type \(V\), and let \(\alpha : V \to \mathrm{Prop}\) be an assignment that makes at least one literal of the clause true. Let \(ml\) be a map lifting each literal over \(V\) to a literal over the extended type \(\mathtt{AuxVar}\, V\), and let \(\alpha _3\) be an assignment on \(\mathtt{AuxVar}\, V\) such that: (i) for every \(k\) the auxiliary variable \(y_{i,k}\) is true under \(\alpha _3\) exactly when all of the first \(k+2\) literals of the clause are false under \(\alpha \); and (ii) for every literal \(\ell \), the lifted literal \(ml(\ell )\) is true under \(\alpha _3\) if and only if \(\ell \) is true under \(\alpha \). Then for every chain counter \(j \ge 1\), letting \(\mathit{lits}\) be the suffix of the clause obtained by dropping its first \(j+1\) literals, every 3-clause in the chain built for clause \(i\) from \(\mathit{lits}\) starting at counter \(j\) is satisfied by \(\alpha _3\).
Fix a clause index \(i\) and an assignment \(\alpha \) of truth values to the auxiliary variables. Let \(\ell \mapsto \widehat{\ell }\) be a map lifting each literal over the base variable type to a literal over the auxiliary-variable type, let \(j \ge 1\), and let \(L\) be a list of at least two literals over the base type such that the lifted literal \(\widehat{\ell }\) evaluates to false under \(\alpha \) for every \(\ell \) in \(L\). If every 3-clause of the chain encoding built from clause index \(i\), the lifting \(\ell \mapsto \widehat{\ell }\), the list \(L\), and the counter \(j\) is satisfied by \(\alpha \), then \(\alpha \) cannot also make the entry variable \(y_{i,\, j-1}\) true; these hypotheses taken together are contradictory.
Let \(i\) be a natural number and let \(ml\) be a lifting map sending each literal over a variable type \(V\) to a literal over \(\mathtt{AuxVar}\, V\), and suppose that \(ml(l)\) is local to index \(i\) for every literal \(l\). Then, for every list of literals \(\mathit{lits}\) over \(V\) and every chain counter \(j\), each \(3\)-clause occurring in the chain of clauses built from \(\mathit{lits}\) (with clause index \(i\), lifting map \(ml\), and counter \(j\)) has all three of its literals local to index \(i\).
Let \(\alpha \) be a truth assignment of the variables \(V\), and let \(f\) be a CNF formula over \(V\) with clauses \(C_0, \dots , C_{m-1}\). Fix an index \(i\) with \(0 \le i \lt m\) and write \(C_i\) for the \(i\)-th clause. Let \(l\) be a literal over the auxiliary variable type \(\mathtt{AuxVar}\, V\) that is local to \(i\), meaning that every auxiliary variable \(y_{i',j}\) occurring in \(l\) satisfies \(i' = i\). Write \(\hat\alpha \) for the global assignment determined by \(\alpha \) and \(f\), which sends each original variable to its value under \(\alpha \) and each auxiliary variable \(y_{i',j}\) to the truth of the assertion that the first \(j+2\) literals of clause \(C_{i'}\) are all false under \(\alpha \); and write \(\alpha _i\) for the corresponding local assignment, built from clause \(C_i\) alone. Then \(l\) takes the same truth value under \(\hat\alpha \) as under \(\alpha _i\).
Given an assignment \(\alpha : V \to \mathrm{Prop}\), SATTo3SAT.evalLiteral maps a literal to its truth value: a positive literal \(\mathtt{pos}\, v\) is true iff \(\alpha (v)\) holds, and a negative literal \(\mathtt{neg}\, v\) is true iff \(\alpha (v)\) does not hold.
For a list of literals \(\mathit{lits}\) and index \(j\), \(\texttt{SATTo3SAT.extraVal}\, \alpha \, \mathit{lits}\, j\) holds iff every literal in \(\mathit{lits}.\mathtt{take}(j + 2)\) is false under \(\alpha \). This is the intended truth value of the auxiliary chain variable \(y_j\) in the encoding of \(\mathit{lits}\).
Fix an assignment \(\alpha \) and a list of literals \(\mathit{lits}\), and let \(j\) be an index with \(j + 2 \le |\mathit{lits}|\). Then the auxiliary variable valuation associated with \(\alpha \), \(\mathit{lits}\), and \(j\) holds if and only if every one of the first \(j + 2\) literals of \(\mathit{lits}\) evaluates to false under \(\alpha \).
Fix an assignment \(\alpha \) of truth values to the variables of a type \(V\), a CNF formula \(f\), and an index \(i\) with \(i \lt \mathrm{length}(f)\); let \(C\) be the \(i\)-th clause of \(f\), and let \(c_3\) be any of the 3-clauses produced by encoding \(C\) as clause number \(i\). Consider the global lifted assignment on \(\mathtt{AuxVar}\, V\), which sends each original variable \(v\) to \(\alpha (v)\) and each auxiliary chain variable \(y_{i',j}\) to the truth value of the assertion “the first \(j+2\) literals of the \(i'\)-th clause of \(f\) are all false under \(\alpha \)”, together with the local assignment that agrees with \(\alpha \) on original variables and sends \(y_{i,j}\) to the truth value of “the first \(j+2\) literals of \(C\) are all false under \(\alpha \)”. Then \(c_3\) is satisfied by the global assignment if and only if it is satisfied by the local assignment.
Given \(\alpha : \mathtt{Assignment}\, V\) and a formula \(f\), SATTo3SAT.globalAssignment \(\alpha \, f\) is the assignment on \(\mathtt{AuxVar}\, V\) that maps each original variable \(\mathtt{orig}\, v\) to \(\alpha (v)\) and each auxiliary variable \(\mathtt{extra}\, i\, j\) to \(\texttt{SATTo3SAT.extraVal}\, \alpha \, (f.\mathtt{get}\, i)\, j\), i.e. the chain invariant “the first \(j+2\) literals of clause \(i\) are all false.”
Let \(l\) be a list over a type \(\alpha \), let \(n\) be a natural number, and suppose that dropping the first \(n\) entries of \(l\) leaves a nonempty list whose head is \(a\) (so that \(l\) has more than \(n\) entries and its \((n+1)\)-st entry is \(a\)). Then the prefix consisting of the first \(n+1\) entries of \(l\) equals the prefix consisting of the first \(n\) entries with \(a\) appended: taking \(n+1\) entries of \(l\) gives (taking \(n\) entries of \(l\)) followed by \([a]\).
Let \(f = (C_0, \dots , C_{m-1})\) be a CNF formula over a variable type \(V\), and write \(\mathrm{Aux}(V)\) for the auxiliary variable type built from \(V\). For each index \(i\) with \(0 \le i \lt m\), let \(T_i\) be the list of \(3\)-clauses obtained by encoding the single clause \(C_i\) at index \(i\), and take the SAT-to-3-SAT encoding of \(f\) to be the \(3\)-CNF formula over \(\mathrm{Aux}(V)\) formed by concatenating \(T_0, T_1, \dots , T_{m-1}\) in order. Then a \(3\)-clause \(c\) over \(\mathrm{Aux}(V)\) occurs in the SAT-to-3-SAT encoding of \(f\) if and only if there exists an index \(i \lt m\) such that \(c\) occurs in \(T_i\).
Let \(\mathit{cs} = (c_0,\dots ,c_{m-1})\) be a list of clauses over a variable type \(V\), let \(\mathit{idx} \in \mathbb {N}\) be a starting index, and let \(c_3\) be a 3-clause over the auxiliary variable type \(\mathtt{AuxVar}\, V\). Encode \(\mathit{cs}\) into 3-clauses by transforming each clause \(c_i\) at the shifted index \(\mathit{idx}+i\) and concatenating the outputs in order of increasing \(i\). Then \(c_3\) occurs in the resulting encoding of \(\mathit{cs}\) if and only if there is a position \(i \lt m\) for which \(c_3\) occurs among the 3-clauses produced by transforming the clause \(c_i\) at index \(\mathit{idx}+i\).
Let \(f\) be a CNF formula over a variable type \(V\), and let \(\mathrm{to3SAT}(f)\) be the 3-CNF formula produced by the standard clause-splitting transformation, whose variables are those of \(V\) together with fresh auxiliary chain variables introduced to break each long clause into a chain of three-literal clauses. If \(f\) is satisfiable — that is, some assignment to \(V\) makes every clause of \(f\) true — then \(\mathrm{to3SAT}(f)\) is 3-satisfiable: there is an assignment to its variables under which every 3-clause is true.
Let \(f\) be a CNF formula over a variable type \(V\), and let \(\tau (f)\) denote its \(3\)-CNF encoding: the formula obtained by replacing each clause of \(f\) with an equisatisfiable chain of \(3\)-clauses over an enlarged variable type that adjoins to \(V\) a supply of fresh auxiliary variables (a clause of more than three literals is split into linked triples using these auxiliary variables, while shorter clauses are padded to three literals). Then \(f\) is satisfiable if and only if \(\tau (f)\) is \(3\)-satisfiable.
Let \(f\) be a CNF formula over a variable type \(V\), and let its SAT-to-3-SAT encoding be the 3-CNF formula over the extended variable type (the original variables together with the fresh auxiliary chain variables). If this encoded 3-CNF formula is 3-satisfiable — that is, some assignment of the extended variables satisfies all of its 3-clauses — then \(f\) itself is satisfiable.
SATTo3SAT.transformClause maps a single clause (at index \(c\_ idx\)) to a list of 3-clauses: the empty clause becomes two contradictory 3-clauses (unsatisfiable), a 1-literal clause becomes \(\langle l_1,l_1,l_1\rangle \), a 2-literal clause becomes \(\langle l_1,l_2,l_2\rangle \), a 3-literal clause is copied as-is, and any longer clause \(l_1 :: l_2 :: \mathit{rest}\) becomes \(\langle l_1, l_2, y_0\rangle \) followed by \(\texttt{SATTo3SAT.buildChain}\) on \(\mathit{rest}\).
Fix a clause index \(i\) and let \(C\) be a clause over a variable type \(V\). If a 3-clause occurs in the list produced by encoding \(C\) at index \(i\), then each of its three literals is local to \(i\); that is, every auxiliary variable appearing in the 3-clause carries the index \(i\), while literals over original variables of \(V\) are local automatically.
Let \(\alpha \) be a truth assignment on a variable set \(V\), and let \(C\) be a clause over \(V\) — a finite list of literals read as their disjunction — carrying the index \(i\). Assume \(\alpha \) satisfies \(C\), meaning at least one literal of \(C\) is true under \(\alpha \). Extend \(\alpha \) to an assignment \(\alpha _3\) on the enlarged variable set consisting of the original variables of \(V\) together with the auxiliary chain variables \(y_{i,j}\): leave each original variable at its value under \(\alpha \), and set \(y_{i,j}\) to true exactly when the first \(j+2\) literals of \(C\) are all false under \(\alpha \). Then every \(3\)-clause in the list obtained by encoding \(C\) at index \(i\) as \(3\)-clauses is satisfied by \(\alpha _3\).
Fix a variable type \(V\), a clause index \(c \in \mathbb {N}\), and a clause \(C\) over \(V\) — a finite list of literals read as their disjunction. Let \(\alpha \) be a truth assignment to the auxiliary variables (the original variables of \(V\) together with the fresh chain variables), and suppose that every \(3\)-clause produced by encoding \(C\) at index \(c\) is satisfied by \(\alpha \). Then \(C\) has at least one literal that is true under the assignment obtained by reading each original variable \(v\) through its copy among the auxiliary variables, i.e. \(v \mapsto \alpha (\mathrm{orig}\, v)\).
Given the Boolean truth values \(a, b, c_3\) of the three literals of a clause, \(\mathrm{clauseGadgetColor}(a, b, c_3, k)\) assigns a color in \(\mathrm{Fin}\, 3\) to gadget node \(k \in \{ 0,\ldots ,5\} \). In particular, node \(5\) always receives color \(1\) (True), and the remaining nodes are colored so that all internal gadget edges receive distinct colors whenever at least one literal is true.
\(\mathrm{EdgeRelation}(f, u, v)\) defines the (undirected) adjacency structure of the reduction graph. The edges are: all pairs of distinct palette nodes; every literal node to the Base palette node; \(\mathrm{pos}(x)\) to \(\mathrm{neg}(x)\) for each variable \(x\); specific pairs of clause gadget nodes within the same clause (encoding two internal triangles on nodes \(\{ 0,1,2\} \) and \(\{ 3,4,5\} \) plus the bridge \(2\)–\(3\)); each clause’s three literal nodes to gadget nodes \(0\), \(1\), and \(4\) respectively; and gadget node \(5\) of every clause to both the Base and the False palette nodes.
Given a type \(V\) of Boolean variables, a literal over \(V\) is either a positive occurrence \(\mathrm{pos}(v)\) or a negative occurrence \(\mathrm{neg}(v)\) of a variable \(v \in V\); a positive literal asserts the variable and a negative literal asserts its negation.
The vertex set of the reduction graph is the inductive type \(\mathrm{OutputVertex}(V)\) with three constructors:
\(\mathrm{palette}(p)\) for \(p : \mathrm{Fin}\, 3\) — the three special palette nodes (Base \(= 0\), True \(= 1\), False \(= 2\)) forming a triangle that fixes color semantics;
\(\mathrm{literalNode}(\ell )\) — one node per literal over \(V\) (positive and negative occurrences are separate nodes);
\(\mathrm{clauseGadget}(c, k)\) for \(k : \mathrm{Fin}\, 6\) — six internal nodes per clause \(c\) encoding the OR constraint.
Given a 3-SAT instance \(f\) over variables \(V\), the reduction graph \(\mathrm{ReductionGraph}(f)\) is the SimpleGraph on \(\mathrm{OutputVertex}(V)\) whose adjacency relation is the symmetrisation of \(\mathrm{EdgeRelation}(f)\): two distinct vertices are adjacent iff \(\mathrm{EdgeRelation}(f, u, v)\) or \(\mathrm{EdgeRelation}(f, v, u)\) holds.
A specific 3-SAT instance over four Boolean variables \(x_0, x_1, x_2, x_3\) (\(V = \mathrm{Fin}\, 4\)), consisting of three clauses: \((x_0 \vee \neg x_1 \vee \neg x_2)\), \((\neg x_0 \vee x_1 \vee \neg x_2)\), and \((\neg x_0 \vee \neg x_1 \vee \neg x_3)\).
Given a satisfying assignment \(\mathrm{assign} : V \to \mathrm{Bool}\), the function \(\mathrm{sat3Coloring}(\mathrm{assign})\) maps every vertex of \(\mathrm{OutputVertex}(V)\) to a color in \(\mathrm{Fin}\, 3\): palette node \(p\) gets color \(p\); a positive (resp. negative) literal node for variable \(v\) gets color \(1\) (True) if \(\mathrm{assign}(v) = \texttt{true}\) and color \(2\) (False) otherwise (resp. the reverse); clause gadget node \((c, k)\) gets \(\mathrm{clauseGadgetColor}\) applied to the truth values of \(c\)’s three literals.
Fix a variable type \(V\) and a Boolean assignment \(\alpha : V \to \{ \text{true}, \text{false}\} \), and consider the coloring of the reduction graph induced by \(\alpha \), which sends each vertex to a color in \(\{ 0, 1, 2\} \) (with \(1\) standing for True and \(2\) for False). For every literal \(\ell \) over \(V\), the literal node of \(\ell \) receives color \(1\) if \(\ell \) is satisfied by \(\alpha \), and color \(2\) otherwise:
Under assignment \(\mathrm{assign}\), a clause \(c\) is satisfied if at least one of its three literals evaluates to true:
Given a Boolean assignment \(\mathrm{assign} : V \to \mathrm{Bool}\), the function \(\texttt{SATtoColor.SatisfiesLiteral}\) evaluates a literal: a positive literal \(\mathrm{pos}(v)\) evaluates to \(\mathrm{assign}(v)\), and a negative literal \(\mathrm{neg}(v)\) evaluates to \(\neg \, \mathrm{assign}(v)\).
An assignment \(\mathrm{assign}\) satisfies a 3-SAT instance \(f\) if every clause in the list \(f\) is satisfied by \(\mathrm{assign}\), i.e. \(\mathrm{SatisfiesSat3}(\mathrm{assign}, f) = \texttt{true}\) iff \(\forall c \in f,\; \mathrm{SatisfiesClause}(\mathrm{assign}, c) = \texttt{true}\).
Let \(f\) be a 3-SAT instance over a variable type \(V\), that is, a list of clauses, each a disjunction of three literals. If \(f\) is satisfiable — there exists a Boolean assignment \(\mathrm{assign} : V \to \mathrm{Bool}\) under which every clause of \(f\) has at least one true literal — then the reduction graph \(\mathrm{ReductionGraph}(f)\) is 3-colorable, i.e. it admits a proper coloring by the three colors \(\mathrm{Fin}\, 3\).
Let \(f\) be a 3-SAT instance over a variable type \(V\): a list of clauses, each clause an ordered triple of literals, where a literal is a positive or a negative occurrence of a variable in \(V\). Then \(f\) is satisfiable — some Boolean assignment \(V \to \{ \mathrm{true}, \mathrm{false}\} \) makes every clause in \(f\) true — if and only if the reduction graph associated to \(f\) is 3-colorable, that is, admits a proper coloring by three colors.
Let \(f\) be a 3-SAT instance over a variable type \(V\), that is, a list of clauses, each clause a triple of literals (positive or negative occurrences of variables). If the reduction graph \(\mathrm{ReductionGraph}(f)\) admits a proper \(3\)-colouring, then \(f\) is satisfiable: there exists a Boolean assignment \(\mathrm{assign} \colon V \to \{ \mathrm{true}, \mathrm{false}\} \) under which every clause of \(f\) has at least one literal evaluating to \(\mathrm{true}\).
Let \(q\) be a prime, let \(G\) be a commutative group, and let \(g \in G\) be an element of order \(q\). For any witness \(w\), randomness \(r\), and challenge \(c\) in \(\mathbb {Z}_q\), the honest transcript — commitment \(a = g^{r}\), challenge \(c\), and response \(s = r + c\, w\) — is accepted by the verifier against the public key \(\mathit{pk} = g^{w}\); that is,
where each exponent is taken via the canonical lift of the corresponding element of \(\mathbb {Z}_q\) to \(\mathbb {N}\).
Let \(G\) be a commutative group, let \(q\) be prime, and let \(g \in G\) be an element of order \(q\), taken as the public generator. Fix a witness \(w \in \mathbb {Z}_q\) and a challenge \(c \in \mathbb {Z}_q\), and let \(\mathit{pk} = g^{w}\) be the associated public key. Then, as functions \(\mathbb {Z}_q \to G \times \mathbb {Z}_q \times \mathbb {Z}_q\) of the prover’s randomness \(r\), the honest transcript map \(r \mapsto \big(g^{r},\, c,\, r + c\cdot w\big)\) coincides with the simulator map fed the reindexed randomness \(\sigma _{w,c}(r) = r + c\cdot w\), namely \(r \mapsto \big(g^{\sigma _{w,c}(r)}\cdot (\mathit{pk}^{\, c})^{-1},\, c,\, \sigma _{w,c}(r)\big)\); that is, these two maps are equal.
Special soundness. Let \(q\) be prime, let \(g\) be an element of order \(q\) in a commutative group \(G\), and take the public key \(\mathit{pk} = g^{w}\) for some witness \(w \in \mathbb {Z}_q\). Suppose \(a \in G\), and suppose \(c_1, c_2 \in \mathbb {Z}_q\) are distinct challenges with responses \(s_1, s_2 \in \mathbb {Z}_q\) such that both transcripts \((a, c_1, s_1)\) and \((a, c_2, s_2)\) are accepted by the verifier for \(\mathit{pk}\). Then the extractor recovers the witness:
The zero-knowledge simulator, given public key \(\mathit{pk}\), challenge \(c\), and a fresh response \(s\), outputs the transcript \((g^s \cdot (\mathit{pk}^c)^{-1},\; c,\; s)\). The commitment is chosen so that the transcript is accepting without using the witness.
A Schnorr transcript is a triple \((a, c, s) \in G \times \mathbb {Z}_q \times \mathbb {Z}_q\), where \(a\) is the prover’s commitment (an element of the group \(G\)), \(c \in \mathbb {Z}_q\) is the verifier’s challenge, and \(s \in \mathbb {Z}_q\) is the prover’s response.
For every natural number \(n \ge 0\), the diagonal Ramsey number satisfies
where \(r(k, l)\) denotes the least number of vertices \(N\) such that every graph on \(N\) vertices contains either a clique on \(k\) vertices or an independent set on \(l\) vertices. (H. L. Abbott)
The associated digraph D(G): each edge becomes two opposite arcs. G.Adj is already symmetric, so it is the arc relation.
The associated digraph D(G) of a graph G is the digraph obtained when each edge e of G is replaced by two oppositely oriented arcs with the same ends as e.
Vertices of attachment of the bridge of e to the cycle traced by c: V(B) \(\cap \) V(C). (B&M’s V(B,H).)
For a bridge B of H, we write V(B) \(\cap \) V(H) = V(B, H), and call the vertices in this set the vertices of attachment of B to H.
In a connected graph every bridge has at least one vertex of attachment, and in a block every bridge has at least two vertices of attachment. A bridge with k vertices of attachment is called a k-bridge. Two k-bridges with the same vertices of attachment are equivalent k-bridges.
Two bridges of a cycle avoid one another: all the attachments of one lie in a single segment of the other.
Two bridges avoid one another if all the vertices of attachment of one bridge lie in a single segment of the other bridge; otherwise they overlap.
Let \(G\) be a finite simple graph on vertex set \(V\), and let \(M\) be a maximum matching of \(G\), so that no matching of \(G\) has more edges than \(M\). For a set \(S \subseteq V\), write \(o(G - S)\) for the number of connected components having an odd number of vertices in the subgraph of \(G\) induced on \(V \setminus S\). Then
where \(\left\lvert M\right\rvert \) is the number of edges of \(M\) and the maximum ranges over all subsets \(S\) of \(V\).
Let \(G\) be a simple graph on a finite vertex set \(V\) that is bipartite with parts \(X\) and \(Y\): the sets \(X\) and \(Y\) are disjoint, together they exhaust \(V\) (so \(X \cup Y = V\)), and every edge of \(G\) joins a vertex of \(X\) to a vertex of \(Y\). For a set \(S \subseteq V\) write \(N(S) = \bigcup _{v \in S} N(v)\) for the set of all vertices adjacent to some vertex of \(S\). Then \(G\) has a perfect matching if and only if \(\left\lvert S\right\rvert \le \left\lvert N(S)\right\rvert \) for every set \(S \subseteq V\).
Let \(G\) be a finite block — a connected graph with no cut vertex — on at least three vertices. Then any two distinct edges \(e_1\) and \(e_2\) of \(G\) lie on a common cycle; that is, there is a cycle in \(G\) whose edge set contains both \(e_1\) and \(e_2\).
Let \(G\) be a finite simple graph on a vertex set with \(n\) vertices, and let \(u\) and \(v\) be two distinct, nonadjacent vertices of \(G\) whose degrees satisfy \(\deg (u) + \deg (v) \ge n\). Then \(G\) is Hamiltonian if and only if the graph \(G + uv\), obtained from \(G\) by adding the edge joining \(u\) and \(v\), is Hamiltonian.
Let \(H\) be a subgraph of a finite simple graph \(G\), and let \(e\) and \(e'\) be edges of \(G\) that do not belong to \(H\), so both lie in \(E(G)\setminus E(H)\). Recall that two such edges are called equivalent when they are joined by a walk, internally disjoint from \(H\), whose first and last edges are the two given ones, and that the bridge of \(H\) containing an edge \(f\) is the subgraph of \(G\) induced by the equivalence class of \(f\). Write \(B\) for the bridge containing \(e\) and \(B'\) for the bridge containing \(e'\). If \(e\) and \(e'\) are not equivalent, then every vertex lying in both \(B\) and \(B'\) is a vertex of \(H\); that is, \(V(B)\cap V(B')\subseteq V(H)\).
Let \(C\) be a cycle in a finite simple graph \(G\), and let \(e\) be an edge of \(G\) that does not lie on \(C\); write \(B\) for the bridge of \(C\) in \(G\) that contains \(e\). Suppose \(v_1, v_2, v_3\) are three distinct vertices of attachment of \(B\), that is, three distinct vertices each lying both in \(B\) and on \(C\). Then there exists a vertex \(v_0\) of \(B\) that does not lie on \(C\), together with three paths \(P_1, P_2, P_3\) in \(G\) joining \(v_0\) to \(v_1\), \(v_2\), and \(v_3\) respectively, such that any two of these paths have only the vertex \(v_0\) in common.
The bridge of H in G containing the edge e: the subgraph induced by e’s ~-class.
A subgraph of G - E(H) induced by an equivalence class under the relation \(\sim \) is called a bridge of H in G. It follows immediately from the definition that if B is a bridge of H, then B is a connected graph and, moreover, that any two vertices of B are connected by a path that is internally-disjoint from H. It is also easy to see that two bridges of H have no vertices in common except, possibly, for vertices of H.
The section then narrows: In this section we are concerned with the study of bridges of a cycle C. Thus, to avoid repetition, we shall abbreviate "bridge of C" to "bridge" in the coming discussion.
Let \(G\) be a finite simple graph, let \(H\) be a subgraph of \(G\), and let \(e\) be an edge of \(G\) that does not belong to \(H\). Then the bridge of \(H\) in \(G\) containing \(e\) — the subgraph induced by the \(\sim \)-equivalence class of \(e\) — is connected: it has at least one vertex, and any two of its vertices are joined by a path within it.
B&M’s relation ~ on E(G) \ E(H): e\(_{1}\) ~ e\(_{2}\) iff joined by a walk internally disjoint from H whose first and last edges are e\(_{1}\), e\(_{2}\). Relation.EqvGen takes the reflexive-transitive-symmetric closure (discharging "~ is an equivalence relation").
Let H be a given subgraph of a graph G. We define a relation \(\sim \) on E(G) \ E(H) by the condition that e\(_{1}\) \(\sim \) e\(_{2}\) if there exists a walk W such that (i) the first and last edges of W are e\(_{1}\) and e\(_{2}\), respectively, and (ii) W is internally-disjoint from H. It is easy to verify that \(\sim \) is an equivalence relation on E(G) \ E(H).
Let \(G\) be a connected simple graph on a finite, nonempty vertex set that is neither a complete graph nor isomorphic to a cycle of odd length. Then the chromatic number of \(G\) is at most its maximum degree, \(\chi (G) \le \Delta (G)\).
The following two statements are equivalent. (i) (Brooks’ theorem) Every finite nonempty connected graph \(G\) that is neither a complete graph nor an odd cycle satisfies \(\chi (G) \le \Delta (G)\), where \(\chi (G)\) denotes the chromatic number and \(\Delta (G)\) the maximum degree. (ii) For every integer \(k \ge 4\), every finite nonempty \(k\)-critical graph \(G\) that is not complete satisfies \(\nu (k-1) + 1 \le 2\varepsilon \), where \(\nu \) is the number of vertices and \(\varepsilon \) the number of edges of \(G\).
Let \(H\) and \(K\) be two simple graphs on a common finite vertex set \(V\), and suppose \(H \le K\), meaning that every edge of \(H\) is also an edge of \(K\). Then \(K\) has no more connected components than \(H\); that is, the number of connected components of \(K\) is at most the number of connected components of \(H\).
Let \(G\) be a simple graph on a finite vertex set \(V\) with \(n = \left\lvert V\right\rvert \) vertices, and let \(G^{c}\) denote its complement, the simple graph on \(V\) in which two distinct vertices are adjacent precisely when they are not adjacent in \(G\). Then the number of edges of \(G\) together with the number of edges of \(G^{c}\) equals the number of unordered pairs of vertices,
Let \(x_1, x_2, \ldots , x_n\) be points in the Euclidean plane \(\mathbb {R}^2\) whose set \(\{ x_1, \ldots , x_n\} \) has diameter exactly \(1\). Then the number of unordered pairs \(\{ x_i, x_j\} \) whose Euclidean distance satisfies \(d(x_i, x_j) \gt 1/\sqrt{2}\) is at most \(\lfloor n^2/3 \rfloor \).
Let \(G\) be a finite simple graph whose chromatic number equals \(k\). Then \(G\) has at least \(k\) vertices of degree at least \(k-1\); that is, the number of vertices \(v\) with \(\deg _G(v) \ge k-1\) is at least \(k\).
For every natural number \(n\), the complete bipartite graph \(K_{n,n}\), whose two vertex classes each have \(n\) vertices and whose edges are exactly the pairs joining a vertex of one class to a vertex of the other, has exactly \(n!\) perfect matchings.
For every natural number \(n\), the number of perfect matchings of the complete graph \(K_{2n}\) on \(2n\) vertices — that is, the number of ways to partition its vertex set into \(n\) disjoint pairs, each joined by an edge — is
For every natural number \(n\), the number of simple graphs on the vertex set \(\{ 1, 2, \ldots , n\} \) equals \(2^{\binom {n}{2}}\), since a simple graph is determined by choosing which of the \(\binom {n}{2}\) possible unordered pairs of distinct vertices form edges.
Let \(x_0, x_1, \dots , x_{n-1}\) be \(n\) points of the Euclidean plane \(\mathbb {R}^2\), and consider the unit-distance graph on them, whose edges join the pairs of points lying at distance exactly \(1\). Then the number of such edges is at most
Let \(x_1, \dots , x_n\) be points of the Euclidean plane \(\mathbb {R}^2\) whose set \(\{ x_1, \dots , x_n\} \) has diameter exactly \(1\). Then the number of unordered pairs \(\{ i, j\} \) with \(i \ne j\) and \(\operatorname{dist}(x_i, x_j) = 1\) is at most \(n\); equivalently, the unit-distance graph on these points has at most \(n\) edges.
For every natural number \(n\), the number of spanning trees of the complete graph \(K_n\) on \(n\) vertices is \(\tau (K_n) = n^{n-2}\), where the exponent is the truncated difference (so it is \(0\) when \(n \lt 2\)).
Let \(G\) be a finite simple graph whose maximum degree \(\Delta (G)\) equals \(3\). Then the line graph of \(G\) — the graph whose vertices are the edges of \(G\), two being adjacent when they share an endpoint — has chromatic number at most \(4\).
Let \(G\) and \(H\) be simple graphs on finite, nonempty vertex sets, and let \(G \vee H\) denote their join, the graph obtained from the disjoint union of \(G\) and \(H\) by adding every edge between a vertex of \(G\) and a vertex of \(H\). Then the chromatic number of the join is the sum of the chromatic numbers of the two factors:
Let \(G\) be a finite simple graph. Suppose that every two odd cycles of \(G\) share a vertex; that is, for any two closed walks in \(G\) that are cycles of odd length, there is a vertex lying on both. Then the chromatic number of \(G\) satisfies \(\chi (G) \le 5\).
Let \(G\) be a finite simple graph on a nonempty vertex set \(V\), with \(\nu = |V|\) vertices and \(\varepsilon \) edges, and let \(\pi _k(G)\) denote the number of proper \(k\)-colourings of \(G\). If \(p \in \mathbb {Z}[x]\) is a polynomial with integer coefficients satisfying \(p(k) = \pi _k(G)\) for every natural number \(k\), then the coefficient of \(x^{\nu - 1}\) in \(p\) equals \(-\varepsilon \).
Let \(G\) be a nonempty finite simple graph on \(\nu \) vertices, and for a natural number \(k\) write \(\pi _k(G)\) for the number of proper \(k\)-colourings of \(G\). Suppose \(p\) is a real polynomial that agrees with this count at every natural number, in the sense that \(p(k) = \pi _k(G)\) for all \(k \in \mathbb {N}\). Then every real root of \(p\) is at most \(\nu \); that is, if \(p(x) = 0\) with \(x \in \mathbb {R}\), then \(x \le \nu \).
Let \(G\) be a finite simple graph on \(n \ge 3\) vertices, and suppose \(G\) is not Hamiltonian. Then there is an integer \(m\) with \(0 \lt m\) and \(2m \lt n\) such that \(G\) is degree-majorised by \(C_{m,n}\) (Chvátal, 1972).
Let \(G\) be a finite simple graph on at least three vertices, and suppose \(G\) is connected. Write \(\alpha (G)\) for the independence number of \(G\) (the largest size of a set of pairwise nonadjacent vertices) and \(\kappa (G)\) for its vertex connectivity. If \(\alpha (G) \le \kappa (G)\), then \(G\) is Hamiltonian; that is, \(G\) has a cycle passing through every vertex exactly once.
Let \(G\) be a finite simple graph on \(n \ge 3\) vertices, and write \(\deg (v)\) for the degree of a vertex \(v\). Suppose that for every integer \(m\) with \(1 \le m\) and \(2m \lt n\), it is not the case that both \(G\) has at least \(m\) vertices of degree at most \(m\) and \(G\) has at least \(n - m\) vertices of degree strictly less than \(n - m\). Then \(G\) is Hamiltonian, that is, \(G\) possesses a Hamiltonian cycle.
Let \(G\) be a simple graph on \(\nu \ge 2\) vertices, with degrees arranged in nondecreasing order as \(d_1 \le d_2 \le \cdots \le d_\nu \). Suppose there is no integer \(m\) with \(1 \le m \lt (\nu +1)/2\) for which both \(d_m \lt m\) and \(d_{\nu -m+1} \lt \nu - m\). Then \(G\) has a Hamilton path, that is, a path that passes through every vertex of \(G\) exactly once. (Chvátal)
Let \(T\) be a tree on \(m\) vertices, and let \(n \ge 1\). Then the generalised Ramsey number \(r(T, K_n)\) — the least integer \(N\) for which every colouring of the edges of the complete graph \(K_N\) with two colours contains either a copy of \(T\) in the first colour or a copy of the complete graph \(K_n\) in the second — is given by
One Bondy–Chvátal closure step: add a nonadjacent pair whose degree sum is \(\ge \) \(\nu \).
Lemma 4.4.1 motivates the following definition. The closure of \(G\) is the graph obtained from \(G\) by recursively joining pairs of nonadjacent vertices whose degree sum is at least \(\nu \) until no such pair remains. We denote the closure of \(G\) by \(c(G)\).
\(c(G)\) is well defined.
Its proof (the "first edge not in G\(_{2}\)" argument) is quoted on hamiltonian_iff_of_closureSteps below, which is where the content is used.
C_{m,n} = K_m \(\lor \) (K_m^c + K_{n-2m}), with + the Mathlib disjoint sum \(\oplus \)g.
Now, for \(1 \le m \lt n/2\), let \(C_{m,n}\) denote the graph \(K_m \vee (K_m^c + K_{n-2m})\), depicted in figure 4.9a; two specific examples, \(C_{1,5}\) and \(C_{2,5}\), are shown in figures 4.9b and 4.9c.
That \(C_{m,n}\) is nonhamiltonian follows immediately from theorem 4.2; for if \(S\) denotes the set of \(m\) vertices of degree \(n-1\) in \(C_{m,n}\), we have \(\omega (C_{m,n} - S) = m + 1 \gt |S|\).
Theorem 4.6 then shows these are the extreme nonhamiltonian graphs: every nonhamiltonian simple graph is degree-majorised by some C_{m,\(\nu \)}. C_{1,\(\nu \)} and C_{2,5} reappear as the extremal cases of Corollary 4.6.
Let \(G\) be a simple graph on a finite, nonempty vertex set. If \(G\) is not connected, then its complement \(G^{c}\) — the simple graph on the same vertices in which two distinct vertices are adjacent precisely when they are non-adjacent in \(G\) — is connected.
Let \(G\) be a simple graph on a finite vertex set, and let \(G^{\mathrm{c}}\) denote its complement, the graph on the same vertices in which two distinct vertices are adjacent exactly when they are non-adjacent in \(G\). If the diameter of \(G\) is greater than \(3\), then the diameter of \(G^{\mathrm{c}}\) is less than \(3\).
For natural numbers \(m\) and \(n\), let \(K_{m,n}\) be the complete bipartite graph whose two parts have \(m\) and \(n\) vertices, with an edge joining every vertex of the first part to every vertex of the second and no edges within either part. Then \(K_{m,n}\) has exactly \(mn\) edges.
For every natural number \(n\), the complete bipartite graph \(K_{n,n}\), whose two vertex classes each have \(n\) vertices and whose edges join every vertex of one class to every vertex of the other, is \(1\)-factorable: its edge set decomposes into edge-disjoint \(1\)-factors (perfect matchings).
For every integer \(n \ge 1\), the complete graph \(K_{2n+1}\) on \(2n+1\) vertices can be expressed as the edge-disjoint union of \(n\) connected \(2\)-factors. Concretely, there exist spanning subgraphs \(H_1, \ldots , H_n\) of \(K_{2n+1}\), each of which is \(2\)-regular (every vertex has exactly two neighbours in it) and connected, whose edge sets are pairwise disjoint and whose union is all of \(K_{2n+1}\).
Let \(G\) be a finite simple graph and let \(e\) be an edge of \(G\). Writing \(c(H)\) for the number of connected components of a graph \(H\), and \(G - e\) for the graph obtained from \(G\) by removing the edge \(e\) (keeping all vertices), one has
That is, deleting a single edge either leaves the number of connected components unchanged or increases it by exactly one.
The composition (lexicographic product) G[H].
7.2.4 The composition of simple graphs \(G\) and \(H\) is the simple graph \(G[H]\) with vertex set \(V(G) \times V(H)\), in which \((u, v)\) is adjacent to \((u', v')\) if and only if either \(uu' \in E(G)\) or \(u = u'\) and \(vv' \in E(H)\).
Let \(G\) be a simple graph on a finite, nonempty vertex set \(V\). Then \(G\) is connected if and only if, for every set \(S \subseteq V\) such that both \(S\) and its complement \(V \setminus S\) are nonempty, there exist vertices \(u \in S\) and \(v \in V \setminus S\) that are adjacent in \(G\).
Let \(G\) be a connected finite simple graph. Then \(G\) is a tree — that is, connected and acyclic — if and only if every edge of \(G\) is a bridge, meaning that deleting it disconnects the graph (leaving its two endpoints in different connected components).
Edge contraction G \(\cdot \) e (§2.4). Here it is stubbed on the same carrier V as a placeholder, purely so that the deletion–contraction recurrence (Theorem 2.8) can be typed.
An edge e of G is said to be contracted if it is deleted and its ends are identified; the resulting graph is denoted G \(\cdot \) e.
G \(\cdot \) uv: contract the edge uv by identifying v into u. Carrier drops to {x // x \(\ne \) v}.
An edge e of G is said to be contracted if it is deleted and its ends are identified; the resulting graph is denoted by G \(\cdot \) e.
Theorem 8.3(ii) restates it locally: G\(_{2}\) \(\cdot \) uv denotes the graph obtained from G\(_{2}\) by identifying u and v.
Let \(G\) be a finite connected graph, let \(T\) be a spanning tree of \(G\), and let \(e\) be an edge of \(T\). Write \(\bar T = E(G) \setminus E(T)\) for the cotree, the set of edges of \(G\) that lie outside \(T\). Then (i) no bond of \(G\) is contained in \(\bar T\), and (ii) there is exactly one bond of \(G\) contained in \(\{ e\} \cup \bar T\).
The covering number \(\beta \)(G). No Nat.sInf \(\emptyset \) = 0 trap: Set.univ always covers, so the set is nonempty.
The number of vertices in a maximum independent set of \(G\) is called the independence number of \(G\) and is denoted by \(\alpha (G)\); similarly, the number of vertices in a minimum covering of \(G\) is the covering number of \(G\) and is denoted by \(\beta (G)\).
The k-cube Q_k.
The \(k\)-cube is the graph whose vertices are the ordered \(k\)-tuples of \(0\)s and \(1\)s, two vertices being joined if and only if they differ in exactly one coordinate.
Defined in chapter 1; chapter 5 uses it only in exercise 5.1.1(a).
The (3,5)-Ramsey graph: ZMod 13, adjacent iff the difference is a cubic residue. {1,5,8,12} is closed under negation, so symm is genuine.
A \((k, l)\)-Ramsey graph is a graph on \(r(k, l) - 1\) vertices that contains neither a clique of \(k\) vertices nor an independent set of \(l\) vertices. [...] Ramsey graphs often seem to possess interesting structures. [...] We get the \((3, 5)\)-Ramsey graph by regarding the thirteen vertices as elements of the field of integers modulo 13, and joining two vertices by an edge if their difference is a cubic residue of 13 (either 1, 5, 8 or 12).
Consider the graph on the thirteen vertices \(\mathbb {Z}/13\mathbb {Z}\) in which two distinct vertices are joined by an edge exactly when their difference is a nonzero cubic residue modulo \(13\), that is, an element of \(\{ 1,5,8,12\} \). This graph contains no clique on three vertices: there is no triple of pairwise adjacent vertices, so the graph is triangle-free.
Let \(G\) be the graph on the thirteen vertices \(\mathbb {Z}/13\mathbb {Z}\) in which two distinct vertices are joined by an edge exactly when their difference is a cubic residue modulo \(13\), that is, an element of \(\{ 1,5,8,12\} \). Then \(G\) contains no independent set of five vertices: among any five vertices, some two are adjacent.
Let \(G\) be a finite simple graph and let \(v,w\) be two of its vertices. The edge \(\{ v,w\} \) is a bridge of \(G\) — that is, its removal increases the number of connected components — if and only if \(v\) and \(w\) are adjacent in \(G\) and the edge \(\{ v,w\} \) occurs in no cycle of \(G\).
Let \(C_5\) denote the cycle graph on five vertices, in which the vertices are arranged in a ring and each vertex is adjacent exactly to its two neighbours. Then \(C_5\) contains no independent set of three vertices; that is, no three of its vertices are pairwise non-adjacent.
Let \(G\) be a simple graph on a finite, nonempty vertex set \(V\), and write \(n = |V|\) for the number of vertices and \(m\) for the number of edges. Let \(\delta (G)\) and \(\Delta (G)\) denote the minimum and maximum degree of \(G\), respectively. Then
G is degree-majorised by H: same order, and every sorted-degree entry of G is \(\le \) the corresponding entry of H.
A sequence of real numbers \((p_1, p_2, \ldots , p_n)\) is said to be majorised by another such sequence \((q_1, q_2, \ldots , q_n)\) if \(p_i \le q_i\) for \(1 \le i \le n\). A graph \(G\) is degree-majorised by a graph \(H\) if \(\nu (G) = \nu (H)\) and the nondecreasing degree sequence of \(G\) is majorised by that of \(H\). For instance, the 5-cycle is degree-majorised by \(K_{2,3}\) because \((2, 2, 2, 2, 2)\) is majorised by \((2, 2, 2, 3, 3)\).
(Erdős, 1970.) Let \(m \ge 1\), and let \(G\) be a finite simple graph that is \(K_{m+1}\)-free, i.e. contains no complete subgraph on \(m+1\) vertices. Then there exists a complete multipartite graph \(H\) on the same number of vertices such that \(G\) is degree-majorised by \(H\); that is, writing the degree sequences of \(G\) and \(H\) in nondecreasing order, each entry of \(G\)’s sequence is at most the corresponding entry of \(H\)’s. Moreover, if \(G\) and \(H\) have the same degree sequence, then \(G\) and \(H\) are isomorphic.
Let \(G\) and \(H\) be simple graphs on a common finite vertex set \(V\). If every vertex has at least as high a degree in \(H\) as in \(G\), that is \(d_G(v) \le d_H(v)\) for all \(v \in V\), then \(G\) is degree-majorised by \(H\).
The sorted (ascending) degree sequence of G.
Listing the vertices as v\(_{1}\), …, v_\(\nu \), the sequence (d(v\(_{1}\)), …, d(v_\(\nu \))) is a degree sequence of G. Chapter 4 always takes it in nondecreasing order — as in the hypothesis of Theorem 4.5 (§4.2, p. 65):
Let \(G\) be a simple graph with degree sequence \((d_1, d_2, \ldots , d_\nu )\), where \(d_1 \leq d_2 \leq \ldots \leq d_\nu \) and \(\nu \geq 3\).
which is why this definition sorts ascending.
Let \(k \ge 2\) be an integer, and let \(G\) be a finite graph on a nonempty vertex set whose chromatic number equals \(k\) and whose girth is at least \(6\). Then there exists a finite graph \(H\) whose chromatic number is at least \(k+1\) and whose girth is at least \(6\).
Let \(G\) be a finite simple graph on \(n\) vertices, where \(n \ge 3\), and suppose every vertex has degree at least \(n/2\) (equivalently, \(n \le 2\delta \), where \(\delta \) denotes the minimum degree of \(G\)). Then \(G\) is Hamiltonian, that is, \(G\) contains a cycle passing through every vertex exactly once.
Let \(m\) and \(n\) be finite index sets, and let \(Q = (Q_{ij})_{i \in m,\, j \in n}\) be a real matrix whose every row sums to \(1\), so that \(\sum _{j \in n} Q_{ij} = 1\) for each \(i \in m\), and whose every column sums to \(1\), so that \(\sum _{i \in m} Q_{ij} = 1\) for each \(j \in n\). Then the two index sets have the same size, \(\left\lvert m\right\rvert = \left\lvert n\right\rvert \); that is, \(Q\) is square.
Let \(G\) be a finite simple graph, and let \(c\) be a closed trail in \(G\) based at a vertex \(u\) — that is, a walk from \(u\) to \(u\) that traverses no edge more than once — of positive length. Then every edge \(e\) appearing in \(c\) lies on a cycle of \(G\): there is a vertex \(v\) and a cycle based at \(v\) whose edge set contains \(e\).
Let \(G\) be a simple graph on a finite vertex set \(V\) with \(n = \left\lvert V\right\rvert \) vertices. Then \(G\) has exactly \(\binom {n}{2}\) edges if and only if \(G\) is the complete graph on \(V\), that is, every pair of distinct vertices is adjacent.
Let \(G\) be a finite simple graph with vertex set \(V\) and edge set \(E\), and let \(\kappa '(G)\) denote its edge connectivity. For every natural number \(k\) with \(k \le \kappa '(G)\),
Let \(G\) and \(H\) be finite simple graphs, and let \(G \square H\) denote their Cartesian product, the graph on \(V(G) \times V(H)\) in which \((u,x)\) and \((v,y)\) are adjacent exactly when either \(u = v\) and \(x \sim y\) in \(H\), or \(x = y\) and \(u \sim v\) in \(G\). For a graph \(K\), write \(\Delta (K)\) for its maximum degree and \(\chi '(K)\) for its chromatic index, the least number of colors needed to color the edges of \(K\) so that any two edges meeting at a common vertex receive different colors (equivalently, the chromatic number of the line graph of \(K\)). If \(H\) has at least one edge, so that \(\Delta (H) \gt 0\), and \(H\) satisfies \(\chi '(H) = \Delta (H)\), then \(\chi '(G \square H) = \Delta (G \square H)\).
Let \(G\) be a finite simple graph, let \(K_2\) denote the complete graph on two vertices, and form the Cartesian product \(G \square K_2\) — the prism over \(G\), whose vertices are the pairs \((v,i)\) with \(v\) a vertex of \(G\) and \(i \in \{ 0,1\} \), and in which \((v,i)\) is adjacent to \((v',i')\) exactly when either \(v = v'\) and \(i \neq i'\), or \(i = i'\) and \(v,v'\) are adjacent in \(G\). Then the edge chromatic number of \(G \square K_2\) — that is, the chromatic number of its line graph — equals its maximum degree \(\Delta (G \square K_2)\).
For every integer \(n \ge 1\), the complete graph \(K_{2n-1}\) on \(2n-1\) vertices has edge-chromatic number equal to \(2n-1\); that is, its edges can be properly coloured with \(2n-1\) colours but no fewer. Equivalently, the chromatic number of the line graph of \(K_{2n-1}\) is \(2n-1\).
Let \(G\) be a finite simple \(k\)-regular graph with \(k \ge 1\) whose vertex set has odd cardinality, and let \(F\) be a set of edges of \(G\) satisfying \(2\left\lvert F\right\rvert \lt k\) (that is, fewer than \(k/2\) edges). Write \(G - F\) for the graph obtained from \(G\) by deleting the edges in \(F\), and let \(\Delta \) denote its maximum degree. Then the chromatic index of \(G - F\)—equivalently, the chromatic number of its line graph—equals \(\Delta + 1\). (Beineke and Wilson.)
Let \(G\) be a finite simple graph whose vertex set has \(2n+1\) elements for some \(n \in \mathbb {N}\), and write \(\Delta \) for its maximum degree and \(\varepsilon \) for its number of edges. If \(\varepsilon \gt n\Delta \), then the edge chromatic number of \(G\) equals \(\Delta + 1\); that is, the least number of colours needed to colour the edges of \(G\) so that no two edges sharing an endpoint receive the same colour is exactly \(\Delta + 1\).
Let \(G\) be a finite simple graph whose number of vertices is odd, and suppose \(G\) is \(k\)-regular for some integer \(k \ge 1\) (every vertex has exactly \(k\) neighbours). Then the edge chromatic number of \(G\) — equivalently, the chromatic number of its line graph — equals \(\Delta (G) + 1\), where \(\Delta (G)\) is the maximum degree of \(G\). Writing \(\chi '(G)\) for the edge chromatic number,
Let \(G\) be a bipartite simple graph on a finite vertex set, and let \(\Delta (G)\) denote its maximum degree. Then the chromatic number of the line graph of \(G\) — equivalently, the least number of colours needed to colour the edges of \(G\) so that no two edges sharing an endpoint receive the same colour — equals \(\Delta (G)\).
Let \(G\) be a \(k\)-regular simple graph, with \(k \ge 2\), on a finite vertex set whose cardinality is even, and let \(uv\) be an edge of \(G\). Form the graph \(G'\) by subdividing the edge \(uv\): delete it and join its ends \(u\) and \(v\) by a path of length two through a single new vertex. Then the chromatic number of the line graph of \(G'\) — equivalently, the edge chromatic number \(\chi '(G')\) — equals \(\Delta (G') + 1\), where \(\Delta (G')\) is the maximum degree of \(G'\).
Edge connectivity \(\kappa \)’(G): the minimum size of an edge cut (sInf \(\emptyset \) = 0).
we then define the edge connectivity \(\kappa '(G)\) of \(G\) to be the minimum \(k\) for which \(G\) has a \(k\)-edge cut. If \(G\) is trivial, \(\kappa '(G)\) is defined to be zero. Thus \(\kappa '(G) = 0\) if \(G\) is either trivial or disconnected, and \(\kappa '(G) = 1\) if \(G\) is a connected graph with a cut edge.
Let \(G\) be a finite simple graph with at least one vertex. Then the edge connectivity of \(G\) is at most its minimum degree; that is, \(\kappa '(G) \le \delta (G)\), where \(\delta (G)\) denotes the least degree of any vertex of \(G\).
The edge covering number \(\beta \)\('\)(G). Nat.sInf \(\emptyset \) = 0 IS a live trap when \(\delta \) = 0 — hence the \(\delta \) > 0 hypothesis in Thm 7.2.
[We denote] the number of edges in a minimum edge covering of \(G\) by \(\beta '(G)\); the numbers \(\alpha '(G)\) and \(\beta '(G)\) are the edge independence number and edge covering number of \(G\), respectively.
The equality subgraph G_l.
If \(l\) is a feasible vertex labelling, we denote by \(E_l\) the set of those edges for which equality holds in (5.11); that is
The spanning subgraph of \(G\) with edge set \(E_l\) is referred to as the equality subgraph corresponding to the feasible vertex labelling \(l\), and is denoted by \(G_l\).
Let \(G\) be a finite simple graph with \(n\) vertices and \(m\) edges, and let \(\delta \) denote its minimum degree. If \(6\delta \le n\) and
then \(G\) is Hamiltonian, that is, \(G\) contains a cycle passing through every vertex exactly once. (P. Erdős)
Let \(G\) be a finite simple graph that is connected and has at least one edge. Then \(G\) is eulerian — that is, it admits an Euler tour, a closed walk traversing each edge exactly once — if and only if every vertex of \(G\) has even degree.
Let \(G\) be a connected simple graph on a finite vertex set \(V\). Then \(G\) admits an Eulerian trail — a walk between some pair of vertices \(u,v\) that traverses every edge of \(G\) exactly once — if and only if the number of vertices of odd degree is at most \(2\).
Let \(G\) be a finite simple graph in which every vertex has even degree. Then the edges of \(G\) can be decomposed into cycles: there is a finite collection of cycles \(C_1, C_2, \ldots , C_m\) in \(G\) such that every edge of \(G\) lies in exactly one of the \(C_i\).
Let \(G\) be a finite bipartite graph with edge set \(E\) of size \(\varepsilon = |E|\) and maximum degree \(\Delta \), and let \(p\) be a positive integer with \(\Delta \le p\). Then there exist \(p\) matchings \(M_1, M_2, \ldots , M_p\) of \(G\) that are pairwise edge-disjoint and satisfy
with each of their sizes controlled by
in particular, any two of the matchings differ in size by at most one.
Let \(G\) be a finite simple graph with vertex set \(V\). Then \(G\) has a spanning subgraph \(H\) that is bipartite (that is, \(2\)-colourable) and in which every vertex keeps at least half of its original degree: writing \(d_G(v)\) for the degree of \(v\) in \(G\) and \(d_H(v)\) for the number of neighbours of \(v\) in \(H\), one has \(d_G(v) \le 2\, d_H(v)\) for every \(v \in V\).
Let \(G\) be a finite connected simple graph that is not Hamiltonian, and let \(C\) be a longest cycle in \(G\) — a cycle whose length is at least that of every cycle of \(G\). Then there is an edge \(e \in E(G) \setminus E(C)\) such that the bridge of \(C\) containing \(e\) has a vertex that does not lie on \(C\).
For every integer \(k \ge 2\), there exists a finite simple graph \(H\) whose chromatic number is at least \(k\) and whose girth is at least \(6\); that is, \(H\) cannot be properly colored with fewer than \(k\) colors, yet every cycle in \(H\) has length at least \(6\).
Let \(G\) be a finite simple graph on a nonempty vertex set, and write \(n\) for the number of its vertices and \(\pi _k(G)\) for the number of distinct proper \(k\)-colourings of \(G\). Then there is a polynomial \(p \in \mathbb {Z}[x]\) with integer coefficients such that \(p(k) = \pi _k(G)\) for every \(k \in \mathbb {N}\); moreover \(p\) has degree \(n\), leading coefficient \(1\), and constant term \(0\), and its coefficients alternate in sign, in the sense that \((-1)^{n-i}\, a_i \ge 0\) for every \(i\), where \(a_i\) denotes the coefficient of \(x^i\) in \(p\).
Let \(H\) be a finite group and let \(K \le H\) be a subgroup. Then there exist a natural number \(n\) and elements \(h_1, \dots , h_n \in H\) such that the assignment \(i \mapsto h_i K\), sending each index to the corresponding left coset viewed as a subset of \(H\), is a bijection from \(\{ 1, \dots , n\} \) onto the collection of all subsets of \(H\), and, simultaneously, the assignment \(i \mapsto K h_i\), sending each index to the corresponding right coset, is also a bijection from \(\{ 1, \dots , n\} \) onto the collection of all subsets of \(H\).
Let \(G\) be a finite simple graph on a nonempty vertex set \(V\), and write \(n\) for the number of vertices and \(m\) for the number of edges of \(G\). If \(n \le m\), then \(G\) contains a cycle: there exist a vertex \(v\) and a closed walk from \(v\) to \(v\) that is a cycle.
Let \(G\) be a finite simple graph whose minimum degree is at least \(2\), meaning every vertex has at least two neighbours. Then \(G\) contains a cycle: there is a vertex \(v\) and a closed walk from \(v\) to itself that is a cycle.
Let \(G\) be a finite bipartite simple graph with minimum degree \(\delta \gt 0\). Then there is an edge colouring \(C\) of \(G\) using \(\delta \) colours — that is, a function assigning to each edge one of the colours \(1, 2, \ldots , \delta \), not required to be proper — such that every colour is represented at every vertex: for each vertex \(v\) and each colour \(i\), some edge incident with \(v\) is assigned colour \(i\).
For every natural number \(n\), there exist points \(x_1, x_2, \ldots , x_n\) in the Euclidean plane \(\mathbb {R}^2\) such that the set \(\{ x_1, x_2, \ldots , x_n\} \) has diameter exactly \(1\) and exactly \(\lfloor n^2/3\rfloor \) of the unordered pairs \(\{ x_i, x_j\} \) satisfy \(d(x_i, x_j) \gt 1/\sqrt{2}\), where \(d\) denotes Euclidean distance.
Let \(G\) be a finite simple graph that is connected but not complete. Then there exist vertices \(u\), \(v\), \(w\) of \(G\) such that \(u\) is adjacent to \(v\) and \(v\) is adjacent to \(w\), yet \(u\) is not adjacent to \(w\).
Let \(G\) be a finite simple graph and let \(k\) be a natural number. If \(G\) is \(k\)-colourable, then \(G\) admits a canonical colouring consisting of at most \(k\) colour classes; that is, there is a list \((V_1, V_2, \dots , V_m)\) of pairwise disjoint sets of vertices with \(m \le k\), whose union is the whole vertex set, in which each \(V_i\) is a maximal independent set of \(G\) among the vertices left uncoloured by \(V_1, \dots , V_{i-1}\).
For every natural number \(n\) with \(n = 4\) or \(n \ge 6\), there is a simple graph \(G\) on \(n\) vertices that is \(4\)-critical: its chromatic number is \(\chi (G) = 4\), and every proper subgraph \(H \subsetneq G\) satisfies \(\chi (H) \lt \chi (G)\).
Let \(G\) be a finite simple graph whose chromatic number is \(\chi (G) = k\). Then \(G\) contains a subgraph \(H\) that is \(k\)-critical, that is, \(H\) satisfies \(\chi (H) = k\) and \(\chi (H') \lt \chi (H)\) for every proper subgraph \(H' \subset H\).
Let \(G\) be a connected simple graph on a finite vertex set \(V\) with at least three vertices, and suppose the edge \(\{ u,v\} \) of \(G\) is a bridge (that is, deleting it disconnects \(G\)). Then \(G\) has a cut vertex: there exists a vertex \(w\) such that the singleton \(\{ w\} \) is a vertex cut of \(G\), meaning \(\{ w\} \) is a proper subset of \(V\) and the graph \(G - w\) obtained by deleting \(w\) is disconnected.
Let \(G\) be a finite bipartite graph with vertex parts \(X\) and \(Y\), so that every edge of \(G\) joins a vertex of \(X\) to a vertex of \(Y\), and suppose \(\left\lvert X\right\rvert = \left\lvert Y\right\rvert = n\). Write \(\varepsilon \) for the number of edges of \(G\), and let \(k\) be a natural number. If \((k-1)\, n \lt \varepsilon \), then \(G\) contains a matching consisting of exactly \(k\) edges.
Let \(G\) be a connected simple graph on a finite vertex set \(V\), and let \(w\) assign a real weight \(w(e)\) to each unordered pair \(e\) of vertices. Then there exists a spanning subgraph \(T\) of \(G\) that is a tree and whose total edge weight \(\sum _{e \in E(T)} w(e)\) is least among all subgraphs \(T'\) of \(G\) that are trees.
Let \(G\) be a finite simple graph, let \(H\) be a subgraph of \(G\), and let \(e\) be an edge of \(G\) that does not belong to \(H\). Then any two vertices of the bridge of \(H\) in \(G\) containing \(e\) are joined by a path in \(G\) that is internally disjoint from \(H\).
Let \(G\) be a finite simple graph with minimum degree \(\delta (G)\), and let \(k\) be a natural number with \(k \le \delta (G)\). Then \(G\) contains a path of length exactly \(k\); that is, there exist vertices \(u\) and \(v\) and a path from \(u\) to \(v\) in \(G\) traversing exactly \(k\) edges.
Let \(G\) be a finite simple graph, and let \(M\) and \(N\) be two disjoint matchings of \(G\) (each a set of edges, no two of which share a vertex) with \(\left\lvert N\right\rvert \lt \left\lvert M\right\rvert \). Then there exist disjoint matchings \(M'\) and \(N'\) of \(G\) such that \(\left\lvert M'\right\rvert = \left\lvert M\right\rvert - 1\), \(\left\lvert N'\right\rvert = \left\lvert N\right\rvert + 1\), and \(M' \cup N' = M \cup N\).
Let \(G\) be a finite bipartite simple graph, and let \(\Delta \) denote its maximum degree. Then there is a finite bipartite simple graph \(H\) that is \(\Delta \)-regular (every vertex of \(H\) has degree exactly \(\Delta \)) together with an embedding of \(G\) into \(H\) — an injection \(V(G) \to V(H)\) under which two vertices of \(G\) are adjacent if and only if their images are adjacent in \(H\). In other words, every finite bipartite graph is (isomorphic to) an induced subgraph of a bipartite graph regular of degree equal to \(\Delta \).
Let \(G\) be a connected simple graph on a finite vertex set \(V\). Then there is a spanning subgraph \(T\) of \(G\) — a simple graph on the same vertex set \(V\) with \(T \le G\) — that is a tree, that is, \(T\) is connected and contains no cycle.
Let \(G\) be a finite simple graph on \(n\) vertices, and suppose that \(n\) is even and that the minimum degree \(\delta \) of \(G\) satisfies \(n + 2 \le 2\delta \), that is, \(\delta \ge \tfrac {n}{2} + 1\). Then \(G\) has a \(3\)-factor: a spanning subgraph in which every vertex has degree exactly \(3\).
Let \(G\) be a finite connected simple graph that is not an odd cycle (that is, \(G\) is not a single cycle of odd length using every edge of \(G\)). Then the edges of \(G\) can be assigned two colours in such a way that at every vertex \(v\) of degree at least \(2\), both colours are represented at \(v\) — each of the two colours appears on some edge incident with \(v\).
Let \(G\) be a finite graph on at least three vertices whose connectivity satisfies \(\kappa (G) \ge 2\). Then for any two distinct vertices \(u\) and \(v\) there exist two \(u\)–\(v\) paths in \(G\) that are internally disjoint and, moreover, share no edge.
The "far pairs" graph: adjacent iff distance exceeds 1/\(\sqrt{}\)2.
Let \(G\) be the graph defined by
and
where \(d(x_i, x_j)\) here denotes the euclidean distance between \(x_i\) and \(x_j\).
Let \(G\) be a finite simple graph whose chromatic number equals \(4\). Then \(G\) contains a subdivision of \(K_4\): there are four distinct branch vertices joined pairwise by six paths that are internally disjoint from one another and internally avoid all four branch vertices.
Let \(G\) be a finite simple graph whose chromatic number equals \(k\). Suppose \(G\) admits a proper colouring, using some number \(n\) of colours, in which every one of the \(n\) colour classes contains at least two vertices. Then \(G\) admits a proper colouring with exactly \(k\) colours in which, likewise, every colour class contains at least two vertices.
Let \(T\) be a tree on \(m\) vertices, and let \(n\) be a natural number for which \(m-1\) divides \(n-1\). Writing \(K_{1,n}\) for the star on \(n+1\) vertices — the complete bipartite graph whose parts have sizes \(1\) and \(n\) — the two-colour generalised Ramsey number of \(T\) against \(K_{1,n}\) satisfies
The generalised Ramsey number r(G\(_{1}\),…,G_m) over Mathlib’s containment \(\sqsubseteq \).
7.2.6 Let \(G_1, G_2, \ldots , G_m\) be simple graphs. The generalised Ramsey number \(r(G_1, G_2, \ldots , G_m)\) is the smallest integer \(n\) such that every \(m\)-edge colouring \((E_1, E_2, \ldots , E_m)\) of \(K_n\) contains, for some \(i\), a subgraph isomorphic to \(G_i\) in colour \(i\).
Let \(G\) be a finite simple graph that is \(k\)-regular, meaning every vertex has exactly \(k\) neighbours, and suppose \(G\) has girth \(5\), so that its shortest cycle has length exactly \(5\). Then \(G\) has at least \(k^2 + 1\) vertices.
Let \(G\) be a finite simple graph that is \(k\)-regular, meaning every vertex has exactly \(k\) neighbors, and suppose \(G\) has girth \(4\), that is, the length of its shortest cycle is \(4\). Then \(G\) has at least \(2k\) vertices.
Let \(G_1\) and \(G_2\) be simple graphs on a common finite, non-empty vertex set, and suppose that each becomes \(k\)-critical once its isolated vertices are discarded, i.e. the subgraph induced on the non-isolated vertices of \(G_i\) is \(k\)-critical for \(i = 1, 2\). Suppose further that these two sets of non-isolated vertices have exactly one vertex \(v\) in common, and let \(vv_1\) be an edge of \(G_1\) and \(vv_2\) an edge of \(G_2\). Then the graph
obtained by deleting the edge \(vv_1\) from \(G_1\) and the edge \(vv_2\) from \(G_2\), forming the union of the two resulting graphs, and adjoining the edge \(v_1v_2\), is again \(k\)-critical.
Let \(G\) be a finite simple graph whose vertex set admits a bipartition into parts \(X\) and \(Y\), so that every edge of \(G\) joins a vertex of \(X\) to a vertex of \(Y\). For a set \(S\) of vertices write \(N(S) = \bigcup _{v \in S} N(v)\) for the union of the neighborhoods of the vertices in \(S\). Then \(G\) has a matching that saturates \(X\) — a set of pairwise disjoint edges covering every vertex of \(X\) — if and only if \(\left\lvert S\right\rvert \le \left\lvert N(S)\right\rvert \) for every \(S \subseteq X\).
Let \(G\) be a finite simple graph on \(\nu \ge 4\) vertices with \(\varepsilon \) edges. If \(G\) is Hamilton-connected — that is, every two distinct vertices of \(G\) are joined by a Hamilton path — then
Let \(G\) be a finite simple graph on \(\nu \) vertices, and suppose \(H\) is obtained from \(G\) by a finite sequence (possibly empty) of Bondy–Chvátal closure steps, each of which selects two distinct nonadjacent vertices whose degrees sum to at least \(\nu \) and joins them by an edge. Then \(G\) is Hamiltonian if and only if \(H\) is Hamiltonian.
Let \(G\) be a finite simple graph on \(\nu \ge 3\) vertices. Suppose the complete graph on the vertices of \(G\) can be obtained from \(G\) by finitely many (possibly zero) Bondy–Chvátal closure steps, each of which joins some pair of distinct nonadjacent vertices whose degrees sum to at least \(\nu \). Then \(G\) is Hamiltonian, that is, \(G\) has a cycle through all of its vertices.
Let \(G\) be a finite simple graph on \(n\) vertices with complement \(G^{c}\), and arrange the vertex degrees of each graph in nondecreasing order, writing \(d_{0} \le d_{1} \le \cdots \le d_{n-1}\) for the sorted degree sequence of \(G\) and \(d'_{0} \le d'_{1} \le \cdots \le d'_{n-1}\) for that of \(G^{c}\). If \(d'_{m} \le d_{m}\) for every index \(m\) with \(2m \le n\), then \(G\) has a Hamiltonian path: there are vertices \(a\) and \(b\) and a walk from \(a\) to \(b\) that visits every vertex of \(G\) exactly once.
Let \(G\) be a finite simple graph, and suppose \(G\) is self-complementary, meaning there is a graph isomorphism between \(G\) and its complement \(G^c\). Then \(G\) has a Hamilton path: there exist vertices \(a\) and \(b\) and a walk from \(a\) to \(b\) in \(G\) that visits every vertex of \(G\) exactly once.
Let \(G\) be a finite simple graph on vertex set \(V\), and suppose \(G\) has a Hamilton path, that is, a path passing through every vertex of \(G\) exactly once. Then for every nonempty proper subset \(S \subsetneq V\), the graph \(G - S\) obtained by deleting the vertices of \(S\) has at most \(\left\lvert S\right\rvert + 1\) connected components.
Let \(G\) be a finite simple graph with vertex set \(V\), and suppose \(G\) has a Hamiltonian cycle. Then for every nonempty proper subset \(S \subsetneq V\), the number of connected components of the induced subgraph \(G - S\) obtained by deleting the vertices of \(S\) satisfies
The Harary graph H_{m,n} on vertex set ZMod n (a circulant).
We shall show that equality holds in (3.1) by constructing an \(m\)-connected graph \(H_{m,n}\) on \(n\) vertices that has exactly \(\{ mn/2\} \) edges. The structure of \(H_{m,n}\) depends on the parities of \(m\) and \(n\); there are three cases. Case 1 \(m\) even. Let \(m = 2r\). Then \(H_{2r,n}\) is constructed as follows. It has vertices \(0, 1, \ldots , n-1\) and two vertices \(i\) and \(j\) are joined if \(i - r \leq j \leq i + r\) (where addition is taken modulo \(n\)). Case 2 \(m\) odd, \(n\) even. Let \(m = 2r + 1\). Then \(H_{2r+1,n}\) is constructed by first drawing \(H_{2r,n}\) and then adding edges joining vertex \(i\) to vertex \(i + (n/2)\) for \(1 \leq i \leq n/2\). Case 3 \(m\) odd, \(n\) odd. Let \(m = 2r + 1\). Then \(H_{2r+1,n}\) is constructed by first drawing \(H_{2r,n}\) and then adding edges joining vertex 0 to vertices \((n-1)/2\) and \((n+1)/2\) and vertex \(i\) to vertex \(i+(n+1)/2\) for \(1 \le i \lt (n-1)/2\).
Let \(m\) and \(n\) be natural numbers with \(n \ge 1\) and \(m \lt n\), and let \(H_{m,n}\) be the Harary graph on the \(n\) vertices \(\mathbb {Z}/n\mathbb {Z}\). Then \(H_{m,n}\) has exactly \(\lceil mn/2 \rceil \) edges (equivalently, \(\lfloor (mn+1)/2 \rfloor \)).
G contains a subdivision of K\(_{4}\): four distinct branch vertices joined pairwise by six paths, internally disjoint from each other and internally avoiding all four branch vertices.
A subdivision of a graph G is a graph that can be obtained from G by a sequence of edge subdivisions. A subdivision of K\(_{4}\) is shown in figure 8.5.
The underlying operation is §3.2: An edge e is said to be subdivided when it is deleted and replaced by a path of length two connecting its ends, the internal vertex of this path being a new vertex.
The book states Hajós’ conjecture (1961) in the same paragraph: if G is k-chromatic, then G contains a subdivision of K_k — adding that the condition is not sufficient; for example, a 4-cycle is a subdivision of K\(_{3}\), but is not 3-chromatic, and that k = 1, 2 are obvious while k = 3 holds because a 3-chromatic graph necessarily contains an odd cycle, and every odd cycle is a subdivision of K\(_{3}\). Theorem 8.5 (Dirac, 1952) settles k = 4.
Let \(G\) be a finite simple graph with \(n \ge 4\) vertices and at least \(2n - 2\) edges. Then \(G\) contains a subdivision of \(K_4\): four distinct branch vertices, joined in pairs by six paths that are internally disjoint from one another and whose interiors avoid all four branch vertices.
Ex 1.2.10: the k-cube — vertices are k-bit strings, adjacent iff they differ in exactly one coordinate.
The k-cube is the graph whose vertices are the ordered k-tuples of 0s and 1s, two vertices being joined if and only if they differ in exactly one coordinate.
For every \(k \in \mathbb {N}\), the \(k\)-cube \(Q_k\) — the graph whose vertices are the length-\(k\) strings of bits, two strings being adjacent precisely when they differ in exactly one coordinate — admits a proper colouring with two colours; that is, \(Q_k\) is bipartite.
Let \(G\) be a simple graph on a finite vertex set \(V\), and write \(n = \left\lvert V\right\rvert \) for its number of vertices. Denote by \(\alpha (G)\) the independence number of \(G\), the largest size of a set of vertices no two of which are adjacent, and by \(\beta (G)\) the covering number of \(G\), the smallest size of a set of vertices meeting every edge. Then
Let \(G\) and \(H\) be simple graphs on finite vertex sets, and let \(G[H]\) denote their composition (lexicographic product): the simple graph with vertex set \(V(G) \times V(H)\) in which \((u,v)\) is adjacent to \((u',v')\) exactly when \(uu' \in E(G)\), or else \(u = u'\) and \(vv' \in E(H)\). Writing \(\alpha (\cdot )\) for the independence number, the independence number of the composition satisfies
Let \(G\) be a finite bipartite graph with vertex classes \(X\) and \(Y\), and suppose that \(G\) has minimum degree \(\delta (G) \gt 0\), so that every vertex is an endpoint of at least one edge. Then the independence number \(\alpha (G)\)—the greatest number of pairwise non-adjacent vertices—equals the edge covering number \(\beta '(G)\), the least number of edges in an edge covering of \(G\) (a set of edges meeting every vertex).
Let \(G\) be a simple graph on a finite vertex set \(V\) that is \(2\)-colourable, meaning it admits a proper colouring with two colours. Then for every subset \(s \subseteq V\), the induced subgraph \(G[s]\) is again \(2\)-colourable.
Let \(V\) be a vertex set and let \(K_V\) denote the complete graph on \(V\), in which every pair of distinct vertices is adjacent. For any subset \(s \subseteq V\), the subgraph of \(K_V\) induced on \(s\) is again a complete graph, namely the complete graph \(K_s\) on the vertex set \(s\): any two distinct vertices of \(s\) are adjacent in the induced subgraph precisely because they are adjacent in \(K_V\).
Two u–v walks are internally disjoint if they share no internal vertex.
A family of paths in \(G\) is said to be internally-disjoint if no vertex of \(G\) is an internal vertex of more than one path of the family.
A finite simple graph \(G\) is bipartite if and only if, for every subgraph \(H\) of \(G\) whose minimum degree \(\delta (H)\) is positive, the independence number \(\alpha (H)\) — the greatest number of pairwise nonadjacent vertices of \(H\) — equals the edge covering number \(\beta '(H)\).
Let \(G\) be a finite simple graph. Then \(G\) is bipartite if and only if every subgraph \(H\) of \(G\) satisfies
where \(V(H)\) is the vertex set of \(H\) and \(\alpha (H)\) denotes the independence number of \(H\), that is, the size of a largest set of pairwise non-adjacent vertices of \(H\).
A block: connected with no cut vertex.
A connected graph that has no cut vertices is called a block. Every block with at least three vertices is 2-connected. A block of a graph is a subgraph that is a block and is maximal with respect to this property. Every graph is the union of its blocks.
A bond (§2.2): a minimal (nonempty) edge cut, i.e. a minimal set of edges whose deletion disconnects G.
A minimal nonempty edge cut of G is called a bond; each cut edge e, for instance, gives rise to a bond {e}. If G is connected, then a bond B of G is a minimal subset of E such that G - B is disconnected.
A canonical colouring, as a list of colour classes.
A k-colouring (V\(_{1}\), V\(_{2}\), …, V_k) of G is said to be canonical if V\(_{1}\) is a maximal independent set of G, V\(_{2}\) is a maximal independent set of G - V\(_{1}\), V\(_{3}\) is a maximal independent set of G - (V\(_{1}\) \(\cup \) V\(_{2}\)), and so on. It is easy to see (exercise 8.6.3) that if G is k-colourable, then there exists a canonical k-colouring of G. By repeatedly using the above method for finding maximal independent sets, one can determine all the canonical colourings of G. The least number of colours used in such a colouring is then the chromatic number of G.
Let \(G\) be a finite simple graph whose vertex set has \(n\) elements, and write \(e(G)\) for its number of edges. Let \(m \ge 1\) be an integer. If
then \(G\) contains the complete bipartite graph \(K_{m,m}\) as a subgraph.
Let \(G\) be a finite simple graph on a vertex set with \(n\) vertices, and let \(m \ge 2\) be an integer. If the number of edges of \(G\) satisfies
then \(G\) contains the complete bipartite graph \(K_{2,m}\) as a subgraph.
Let \(G\) be a finite simple graph with vertex set \(V\), write \(n = \left\lvert V\right\rvert \) for its number of vertices and \(d(v)\) for the degree of a vertex \(v\), and let \(m \ge 2\) be an integer. If
then \(G\) contains a copy of the complete bipartite graph \(K_{2,m}\).
G is critical: \(\chi \)(H) < \(\chi \)(G) for every proper subgraph H \(\subset \) G.
It is helpful, when dealing with colourings, to study the properties of a special class of graphs called critical graphs. We say that a graph G is critical if \(\chi \)(H) < \(\chi \)(G) for every proper subgraph H of G. Such graphs were first investigated by Dirac (1952).
Let \(G\) be a critical graph on a finite, nonempty vertex set, so that \(\chi (H) \lt \chi (G)\) for every proper subgraph \(H\) of \(G\). Then for any two distinct vertices \(u\) and \(v\), the neighbourhood \(N(u)\) is not contained in the neighbourhood \(N(v)\).
Let \(G\) be a finite simple graph that is critical, meaning that every proper subgraph of \(G\) has strictly smaller chromatic number than \(G\). Suppose \(u\) and \(v\) are two distinct vertices such that \(\{ u,v\} \) is a vertex cut of \(G\), that is, deleting \(u\) and \(v\) leaves the remaining graph disconnected. Then \(u\) and \(v\) are not adjacent in \(G\).
Let \(G\) be a finite connected graph that is critical, meaning that every proper subgraph \(H\) of \(G\) satisfies \(\chi (H) \lt \chi (G)\), where \(\chi \) denotes the chromatic number. Then no vertex cut of \(G\) is a clique: if \(S\) is a proper set of vertices whose deletion leaves \(G\) disconnected, then the vertices of \(S\) are not pairwise adjacent.
v is a cut vertex: G is connected but deleting v disconnects it.
A vertex \(v\) of \(G\) is a cut vertex if \(E\) can be partitioned into two nonempty subsets \(E_1\) and \(E_2\) such that \(G[E_1]\) and \(G[E_2]\) have just the vertex \(v\) in common. If \(G\) is loopless and nontrivial, then \(v\) is a cut vertex of \(G\) if and only if \(\omega (G - v) \gt \omega (G)\).
An edge covering L: every vertex is an end of some edge of L.
The edge analogue of a covering is called an edge covering. An edge covering of \(G\) is a subset \(L\) of \(E\) such that each vertex of \(G\) is an end of some edge in \(L\). Note that edge coverings do not always exist; a graph \(G\) has an edge covering if and only if \(\delta \gt 0\).
An edge cut: a set of edges of G whose deletion leaves G disconnected.
Recall that an edge cut of \(G\) is a subset of \(E\) of the form \([S, \bar{S}]\), where \(S\) is a nonempty proper subset of \(V\). A \(k\)-edge cut is an edge cut of \(k\) elements. If \(G\) is nontrivial and \(E'\) is an edge cut of \(G\), then \(G - E'\) is disconnected.
An edge cut set: a set of edges whose deletion disconnects G.
For subsets S and S’ of V, [S, S’] denotes the set of edges with one end in S and the other in S’. An edge cut of G is a subset of E of the form [S, \(\bar{S}\)], where S is a nonempty proper subset of V and \(\bar{S}\) = V \ S.
G is eulerian: it has a closed Euler trail (an Euler tour).
A trail that traverses every edge of \(G\) is called an Euler trail of \(G\) because Euler was the first to investigate the existence of such trails in graphs.
A tour of \(G\) is a closed walk that traverses each edge of \(G\) at least once. An Euler tour is a tour which traverses each edge exactly once (in other words, a closed Euler trail). A graph is eulerian if it contains an Euler tour.
Feasible vertex labelling: l x + l y \(\ge \) w s(x,y).
We define a feasible vertex labelling as a real-valued function \(l\) on the vertex set \(X \cup Y\) such that, for all \(x \in X\) and \(y \in Y\)
(The real number \(l(v)\) is called the label of the vertex \(v\).) [...] No matter what the edge weights are, there always exists a feasible vertex labelling; one such is the function \(l\) given by
G is Hamilton-connected: every ordered pair is joined by a Hamilton path (Ex 4.2.11).
4.2.11 \(G\) is Hamilton-connected if every two vertices of \(G\) are connected by a Hamilton path. (a) Show that if \(G\) is Hamilton-connected and \(\nu \ge 4\), then \(\varepsilon \ge [\frac{1}{2}(3\nu +1)]\). (b)* For \(\nu \ge 4\), construct a Hamilton-connected graph \(G\) with \(\varepsilon = [\frac{1}{2}(3\nu +1)]\). (J. W. Moon)
Let \(G\) be a finite simple graph that is \(3\)-regular and admits a proper \(3\)-edge-colouring. If \(G\) is uniquely \(3\)-edge-colourable — meaning any two proper \(3\)-edge-colourings of \(G\) agree up to a permutation of the three colours — then \(G\) contains a Hamiltonian cycle, that is, a closed walk that visits every vertex exactly once.
G is hypohamiltonian: not Hamiltonian, but every vertex-deleted subgraph is (Ex 4.2.12).
4.2.12 \(G\) is hypohamiltonian if \(G\) is not hamiltonian but \(G - v\) is hamiltonian for every \(v \in V\). Show that the Petersen graph (figure 4.4) is hypohamiltonian. (Herz, Duby and Vigué, 1967 have shown that it is, in fact, the smallest such graph.)
Let \(G\) be a graph with vertex set \(V\), and let \(S \subseteq V\). Then \(S\) is an independent set of \(G\) — no edge of \(G\) has both of its ends in \(S\) — if and only if the complement \(V \setminus S\) is a covering of \(G\), that is, every edge of \(G\) has at least one end in \(V \setminus S\).
G is k-critical: k-chromatic and critical.
A k-critical graph is one that is k-chromatic and critical; every k-chromatic graph has a k-critical subgraph. A 4-critical graph, due to Grötzsch (1958), is shown in figure 8.2.
Let \(G\) be a finite graph on a nonempty vertex set, and let \(k\) be a natural number. If \(G\) is \(k\)-critical — that is, \(G\) has chromatic number \(k\) and every proper subgraph of \(G\) has strictly smaller chromatic number — then the edge connectivity of \(G\) satisfies \(\kappa '(G) \ge k - 1\); equivalently, \(G\) is \((k-1)\)-edge-connected.
Let \(G\) be a finite \(k\)-critical graph, that is, a graph with chromatic number \(\chi (G) = k\) every proper subgraph of which has strictly smaller chromatic number. Then every vertex of \(G\) has degree at least \(k - 1\); equivalently, the minimum degree satisfies \(\delta (G) \ge k - 1\).
Let \(G\) be a finite simple graph on a nonempty vertex set, and let \(k\) be a natural number such that \(G\) is \(k\)-critical (that is, \(\chi (G) = k\) and \(\chi (H) \lt k\) for every proper subgraph \(H\) of \(G\)). Let \(S\) be a vertex cut of \(G\), i.e. a proper subset of the vertices whose deletion leaves \(G\) disconnected. Then the subgraph of \(G\) induced on \(S\) is not uniquely \((k-1)\)-colourable: there exist two colourings of this induced subgraph with \(k-1\) colours that induce different partitions of \(S\) into colour classes.
G is k-factorable: it decomposes into edge-disjoint k-factors.
\(G\) is \(k\)-factorable if there are edge-disjoint \(k\)-factors \(H_1, H_2, \ldots , H_n\) such that \(G = H_1 \cup H_2 \cup \ldots \cup H_n\).
Let \(\mathscr {C}\) be an optimal \(k\)-edge colouring of a finite graph \(G\), and suppose there is a vertex \(u\) and colours \(i, j\) for which colour \(i\) is not represented at \(u\) while at least two edges incident with \(u\) carry colour \(j\). Then the connected component of \(u\) in the subgraph \(G[E_i \cup E_j]\) spanned by the edges coloured \(i\) or \(j\) is an odd cycle: there is a closed walk based at \(u\) that is a cycle of odd length whose vertices are precisely those of that component.
An optimal k-edge colouring: one that cannot be improved, where an improvement strictly increases \(\sum \) v, c(v). A static extremal condition, not a procedure.
We shall call a \(k\)-edge colouring \(\mathscr {C}'\) an improvement on \(\mathscr {C}\) if
where \(c'(v)\) is the number of distinct colours represented at \(v\) in the colouring \(\mathscr {C}'\). An optimal \(k\)-edge colouring is one which cannot be improved.
An optimal matching: a max-weight perfect matching.
The optimal assignment problem is clearly equivalent to that of finding a maximum-weight perfect matching in this weighted graph. We shall refer to such a matching as an optimal matching.
Let \(G\) be a finite graph and let \(C\) be a \(k\)-edge colouring of \(G\), that is, an assignment of one of the colours \(1, 2, \ldots , k\) to each edge. Then \(C\) is proper — no two distinct edges that share a common endpoint receive the same colour — if and only if \(c(v) = d(v)\) for every vertex \(v\), where \(c(v)\) is the number of distinct colours represented at \(v\) and \(d(v)\) is the degree of \(v\).
Fix natural numbers \(k\) and \(l\), and say that \(n\) is a Ramsey bound for \((k,l)\) if every simple graph on \(n\) vertices contains either a clique on \(k\) vertices or an independent set on \(l\) vertices. If \(n \le m\) and \(n\) is a Ramsey bound for \((k,l)\), then \(m\) is also a Ramsey bound for \((k,l)\).
For natural numbers \(k\) and \(l\), let \(r(k,l)\) denote the Ramsey number, defined as the least \(n\) (equivalently, the infimum over all such \(n\)) for which every simple graph on \(n\) vertices contains either a clique of size \(k\) or an independent set of size \(l\). Then \(r(k,l)\) itself has this property: every simple graph on \(r(k,l)\) vertices contains a clique of size \(k\) or an independent set of size \(l\).
IsRamseyBoundMulti n c: every m-edge-colouring of K\(_{n}\) has, for some i, a set of c i vertices whose internal edges are all colour i.
We define \(r(k_1, k_2, \ldots , k_m)\) to be the smallest integer \(n\) such that every \(m\)-edge colouring \((E_1, E_2, \ldots , E_m)\) of \(K_n\) contains, for some \(i\), a complete subgraph on \(k_i\) vertices, all of whose edges are in colour \(i\).
Colour i is represented at v: some edge incident with v has colour i. A k-edge colouring here is a plain function C : G.edgeSet \(\to \) Fin k — deliberately not proper.
A \(k\)-edge colouring \(\mathscr {C}\) of a loopless graph \(G\) is an assignment of \(k\) colours, \(1, 2, \ldots , k\), to the edges of \(G\). The colouring \(\mathscr {C}\) is proper if no two adjacent edges have the same colour.
We say that colour \(i\) is represented at vertex \(v\) if some edge incident with \(v\) has colour \(i\).
Let \(G\) be a finite tree, and for a graph \(H\) write \(o(H)\) for the number of connected components of \(H\) having an odd number of vertices. Then \(G\) has a perfect matching if and only if \(o(G - v) = 1\) for every vertex \(v\) of \(G\), where \(G - v\) denotes the graph obtained from \(G\) by deleting \(v\) together with all edges incident to it.
G is uniquely k-edge-colourable: any two proper k-edge colourings agree up to a permutation of the colours.
6.2.5 \(G\) is called uniquely \(k\)-edge-colourable if any two proper \(k\)-edge colourings of \(G\) induce the same partition of \(E\).
A vertex cut: a proper subset S whose deletion leaves G disconnected.
A vertex cut of \(G\) is a subset \(V'\) of \(V\) such that \(G - V'\) is disconnected. A \(k\)-vertex cut is a vertex cut of \(k\) elements. A complete graph has no vertex cut; in fact, the only graphs which do not have vertex cuts are those that contain complete graphs as spanning subgraphs.
Join G \(\lor \) H, rebuilt on Mathlib: within-side edges from \(\oplus \)g, all cross edges from completeBipartiteGraph. symm/loopless are inherited from \(\sqcup \).
We first introduce the notion of the join of two graphs. The join \(G \vee H\) of disjoint graphs \(G\) and \(H\) is the graph obtained from \(G + H\) by joining each vertex of \(G\) to each vertex of \(H\); it is represented diagrammatically as in figure 4.8.
Let \(G\) and \(H\) be simple graphs on finite, nonempty, disjoint vertex sets, and let \(G \vee H\) denote their join, the graph obtained from the disjoint union of \(G\) and \(H\) by adding an edge between every vertex of \(G\) and every vertex of \(H\). Recall that a graph is critical when every proper subgraph has strictly smaller chromatic number than the graph itself. Then \(G \vee H\) is critical if and only if both \(G\) and \(H\) are critical.
Let \(G\) be a finite graph, let \(X\) be a set of its vertices with complement \(V \setminus X\), and let \(n\) be a natural number. Suppose that the subgraphs induced on \(X\) and on \(V \setminus X\) are each \(n\)-colourable, and that the edge cut between the two parts contains fewer than \(n\) edges — that is, the number of edges having one endpoint in \(X\) and the other in \(V \setminus X\) is at most \(n-1\). Then \(G\) is itself \(n\)-colourable.
Let \(G\) be a finite, nonempty \(k\)-critical graph, and let \(u,v\) be two distinct vertices such that \(\{ u,v\} \) is a vertex cut of \(G\), meaning that deleting \(u\) and \(v\) leaves the remaining graph disconnected. Then the degrees of \(u\) and \(v\) satisfy
equivalently \(d(u) + d(v) \ge 3k - 5\).
(Dirac, 1953.) Let \(G\) be a finite, nonempty \(k\)-critical graph, and let \(u \neq v\) be two vertices for which \(\{ u,v\} \) is a \(2\)-vertex cut of \(G\). Then \(G - \{ u,v\} \) has two distinct components \(C_1\) and \(C_2\) whose \(\{ u,v\} \)-components \(G_1 = G[C_1 \cup \{ u,v\} ]\) and \(G_2 = G[C_2 \cup \{ u,v\} ]\) together make up all of \(G\), that is \(G = G_1 \cup G_2\); here \(G_1\) is of type \(1\) (every \((k-1)\)-colouring of \(G_1\) assigns \(u\) and \(v\) the same colour) and \(G_2\) is of type \(2\) (every \((k-1)\)-colouring of \(G_2\) assigns \(u\) and \(v\) different colours). Moreover, the graph \(G_1 + uv\) obtained by adding the edge \(uv\) to \(G_1\) and the graph \(G_2 \cdot uv\) obtained by identifying \(u\) and \(v\) in \(G_2\) are both \(k\)-critical.
Let \(G\) be a finite bipartite graph with bipartition \((X, Y)\) of its vertex set. If \(M\) is a maximum matching of \(G\) and \(K\) is a minimum covering of \(G\), then the number of edges in \(M\) equals the number of vertices in \(K\).
Let \(A\) be an \(m \times n\) matrix with entries in \(\{ 0,1\} \), and call a *line* of \(A\) any single row or single column. Consider two quantities: the minimum number \(c\) of lines needed to cover all the \(1\)-entries of \(A\) (that is, the least value of \(\left\lvert R\right\rvert + \left\lvert C\right\rvert \) over all sets \(R\) of rows and \(C\) of columns such that every entry equal to \(1\) lies in a row of \(R\) or a column of \(C\)), and the maximum number \(d\) of \(1\)-entries of \(A\) no two of which share a line (that is, the largest size of a set of positions with entry \(1\) whose members have pairwise distinct rows and pairwise distinct columns). Then \(c = d\).
Let \(G\) be a finite bipartite graph with bipartition \((X, Y)\), and for a set \(S \subseteq X\) write \(N(S) = \bigcup _{v \in S} N(v)\) for the set of vertices adjacent to some vertex of \(S\). If \(M\) is a maximum matching of \(G\), then
where \(\lvert M\rvert \) denotes the number of edges of \(M\).
Let \(G\) be a finite simple graph that is \(3\)-regular, so every vertex has exactly three neighbours, and suppose \(G\) admits a Hamiltonian cycle, a closed walk that visits every vertex exactly once. Then the line graph of \(G\) — whose vertices are the edges of \(G\), two being adjacent precisely when they share an endpoint — is \(3\)-colourable; equivalently, the edges of \(G\) can be properly coloured with three colours, a Tait colouring.
Let \(G\) be a finite simple graph, and let \(L(G)\) denote its line graph, whose vertices are the edges of \(G\), with two such vertices adjacent precisely when the corresponding edges of \(G\) share a common endpoint. Then the number of edges of \(L(G)\) equals
where the sum runs over all vertices \(v\) of \(G\) and \(\deg _G(v)\) is the degree of \(v\).
Let \(G\) be a connected simple graph on a finite vertex set, with minimum degree \(\delta (G)\) and order \(n\). If \(2\, \delta (G) \lt n\), then \(G\) contains a path of length at least \(2\, \delta (G)\); that is, there exist vertices \(a\) and \(b\) and a path from \(a\) to \(b\) in \(G\) with at least \(2\, \delta (G)\) edges.
Let \(G\) be a simple graph on a finite vertex set that is bipartite with parts \(X\) and \(Y\), so that every edge of \(G\) joins a vertex of \(X\) to a vertex of \(Y\). If there is an integer \(k \gt 0\) such that every vertex of \(G\) has degree exactly \(k\), then \(G\) admits a perfect matching: a set of edges that is vertex-disjoint and covers every vertex of \(G\).
Let \(G\) be a finite simple graph. If \(M\) is a matching in \(G\) and \(K\) is a covering of \(G\), then the number of edges of \(M\) is at most the number of vertices in \(K\); that is, \(\left\lvert M\right\rvert \le \left\lvert K\right\rvert \).
The matching number \(\alpha \)\('\)(G).
The edge analogue of an independent set is a set of links no two of which are adjacent, that is, a matching. [...] We denote the number of edges in a maximum matching of \(G\) by \(\alpha '(G)\) [...] the edge independence number.
Let \(G\) be a finite simple graph on \(n\) vertices with minimum degree \(\delta (G) \gt 0\) (so that \(G\) has no isolated vertices, and edge coverings exist). Then the matching number and the edge covering number of \(G\) satisfy
where \(\alpha '(G)\) is the largest number of edges in a matching of \(G\) and \(\beta '(G)\) is the smallest number of edges in an edge covering of \(G\).
This is the theorem of Gallai (1959).
Let \(G\) be a finite simple graph, and let \(L(G)\) denote its line graph — the graph whose vertices are the edges of \(G\), with two such vertices adjacent exactly when the corresponding edges of \(G\) share an endpoint. Then the maximum degree \(\Delta (G)\) of \(G\) is at most the chromatic number of \(L(G)\); that is, writing \(\chi '(G) = \chi (L(G))\) for the edge chromatic number of \(G\), one has
Let \(G\) be a simple graph on a finite vertex set with \(n \ge 3\) vertices, and write \(E\) for its number of edges. If \(E + 6 \le 3n\) — equivalently \(E \le 3n - 6\) — then \(G\) has a vertex of degree at most \(5\); that is, the minimum degree of \(G\) satisfies \(\delta (G) \le 5\).
Mycielski’s construction on V \(\oplus \) V \(\oplus \) Unit: inl = the old graph, inr \(\circ \) inl = the shadows u\(_{i}\) (joined to v\(_{i}\)’s neighbours), inr \(\circ \) inr = the apex (joined to every shadow).
Suppose that we have already constructed a triangle-free graph G_k with chromatic number k \(\ge \) 2. Let the vertices of G_k be v\(_{1}\), v\(_{2}\), …, v_n. Form a new graph G_{k+1} from G_k as follows: add n+1 new vertices u\(_{1}\), u\(_{2}\), …, u_n, v, and then, for 1 \(\le \) i \(\le \) n, join u\(_{i}\) to the neighbours of v\(_{i}\) and to v. For example, if G\(_{2}\) is K\(_{2}\) then G\(_{3}\) is the 5-cycle and G\(_{4}\) the Grötzsch graph (see figure 8.10).
For every integer \(k \ge 2\), the graph \(G_k\) in the Mycielski tower is \(k\)-critical: its chromatic number is \(\chi (G_k) = k\), and it is critical, meaning that every proper subgraph \(H \subsetneq G_k\) satisfies \(\chi (H) \lt \chi (G_k)\).
Let \(k \ge 1\) be an integer, and let \(G\) be a finite simple graph on \(4k+1\) vertices in which every vertex has degree exactly \(2k\). Then \(G\) has a Hamiltonian cycle, that is, a cycle passing through every vertex of \(G\) exactly once.
Let \(G\) be a finite simple graph, and write \(\alpha (G)\) for its independence number (the size of a largest set of pairwise non-adjacent vertices). Suppose \(G\) is connected and \(\alpha \)-critical, meaning that \(\alpha (G - e) \gt \alpha (G)\) for every edge \(e\) of \(G\). Then \(G\) has no cut vertex: for every vertex \(v\), the singleton \(\{ v\} \) is not a vertex cut, that is, it is not the case that \(\{ v\} \) is a proper subset of the vertex set whose deletion leaves the graph induced on the remaining vertices disconnected.
Let \(G\) be a finite connected simple graph that is \(\beta \)-critical, meaning that deleting any single edge strictly decreases its covering number \(\beta (G)\) (the least number of vertices meeting every edge). Then no single vertex is a cut vertex: for every vertex \(v\), the singleton \(\{ v\} \) is not a vertex cut, that is, removing \(v\) never leaves the graph disconnected.
Let \(G\) be a finite graph on a nonempty vertex set \(V\). If \(G\) is \(k\)-critical — that is, \(G\) has chromatic number \(\chi (G) = k\) and every proper subgraph \(H\) of \(G\) satisfies \(\chi (H) \lt \chi (G)\) — then \(G\) does not have exactly \(k+1\) vertices; that is, \(\left\lvert V\right\rvert \neq k+1\).
Let \(G\) be a simple graph on a finite vertex set \(V\) and \(H\) a simple graph on a vertex set \(W\). Then \(G\) and \(H\) are isomorphic if and only if there is a bijection \(\theta \colon V \to W\) such that for all vertices \(u, v \in V\), the pair \(u, v\) is adjacent in \(G\) if and only if \(\theta (u), \theta (v)\) is adjacent in \(H\).
Let \(G\) be a finite simple graph on at least three vertices, and suppose its vertex connectivity satisfies \(\kappa (G) \lt 2\) (equivalently, \(G\) is not \(2\)-connected). Then \(G\) is nonhamiltonian; that is, \(G\) has no Hamiltonian cycle.
Let \(G\) be a finite simple graph, and let \(X\) and \(Y\) be disjoint sets of its vertices such that every edge of \(G\) joins a vertex of \(X\) to a vertex of \(Y\). If \(\left\lvert X\right\rvert \neq \left\lvert Y\right\rvert \), then \(G\) is not Hamiltonian; that is, \(G\) has no cycle passing through every vertex of \(G\).
Let \(G\) be a simple graph on a finite vertex set \(V\) with \(n = \left\lvert V\right\rvert \) vertices, and let \(G^{c}\) denote its complement, the graph on \(V\) in which two distinct vertices are adjacent precisely when they are non-adjacent in \(G\). Writing \(\chi \) for the chromatic number, the sum of the chromatic numbers of a graph and its complement satisfies
Let \(G\) be a simple graph on a finite vertex set \(V\) with \(n = \left\lvert V\right\rvert \ge 11\), write \(E(G)\) for its edge set, and let \(G^{c}\) be its complement (the simple graph on \(V\) in which two distinct vertices are adjacent exactly when they are not adjacent in \(G\)). Then \(G\) and \(G^{c}\) cannot both satisfy \(\left\lvert E(G)\right\rvert + 6 \le 3n\) and \(\left\lvert E(G^{c})\right\rvert + 6 \le 3n\); that is, at least one of the two graphs has more than \(3n - 6\) edges.
\(\pi \)_k(G), the number of distinct proper k-colourings.
We shall denote the number of distinct k-colourings of G by \(\pi \)_k(G); thus \(\pi \)_k(G) > 0 if and only if G is k-colourable. Two colourings are to be regarded as distinct if some vertex is assigned different colours in the two colourings; in other words, if (V\(_{1}\), V\(_{2}\), …, V_k) and (V\(_{1}\)’, V\(_{2}\)’, …, V_k’) are two colourings, then (V\(_{1}\), V\(_{2}\), …, V_k) = (V\(_{1}\)’, V\(_{2}\)’, …, V_k’) if and only if V\(_{i}\) = V\(_{i}\)’ for 1 \(\le \) i \(\le \) k. For example, a triangle has the six distinct 3-colourings shown in figure 8.7. Note that even though there is exactly one vertex of each colour in each colouring, we still regard these six colourings as distinct.
The book then records the two anchoring cases: If G is empty, then each vertex can be independently assigned any one of the k available colours. Therefore \(\pi \)_k(G) = k^\(\nu \). On the other hand, if G is complete, then there are k choices of colour for the first vertex, k-1 choices for the second, k-2 for the third, and so on. Thus, in this case, \(\pi \)_k(G) = k(k-1)…(k-\(\nu \)+1).
Let \(G\) be a finite simple graph, let \(u\) and \(v\) be adjacent vertices of \(G\), and let \(k\) be a natural number. Writing \(\pi _k(H)\) for the number of proper \(k\)-colourings of a graph \(H\), let \(G - uv\) be the graph obtained from \(G\) by deleting the edge \(uv\), and let \(G \cdot uv\) be the graph obtained by contracting that edge, i.e. by identifying \(u\) and \(v\) into a single vertex. Then
Let \(G\) be a finite simple graph and \(k\) a natural number, and write \(\pi _k(G)\) for the number of proper \(k\)-colourings of \(G\). Then
where the product ranges over the connected components \(C\) of \(G\) and \(G[C]\) denotes the subgraph of \(G\) induced on the vertex set of \(C\).
Let \(G\) be a finite graph, and let \(\pi _k(H)\) denote the number of distinct proper \(k\)-colourings of a graph \(H\). Write \(G \vee K_1\) for the join of \(G\) with a single vertex — the graph obtained from \(G\) by adjoining one new vertex and joining it to every vertex of \(G\). Then for every integer \(k \ge 1\),
Let \(G\) be a connected simple graph on a finite, nonempty vertex set of size \(\nu \), and let \(k\) be a natural number. Writing \(\pi _k(G)\) for the number of distinct proper \(k\)-colourings of \(G\), one has
and moreover, if \(k \ge 3\), then equality \(\pi _k(G) = k(k-1)^{\nu -1}\) holds if and only if \(G\) is a tree.
Let \(G\) be a simple graph on a finite vertex set, and let \(G'\) and \(H'\) be two subgraphs of \(G\); write \(\pi _k(\cdot )\) for the number of distinct proper colourings of a graph with \(k\) available colours, and regard each of \(G'\), \(H'\), their join \(G' \cup H'\), and their meet \(G' \cap H'\) as a graph on its own vertex set. Suppose the intersection \(G' \cap H'\) is complete, that is, every two of its vertices are adjacent in it. Then for every \(k \in \mathbb {N}\),
where \(G' \cup H'\) is written for the join.
Let \(n \ge 3\) and \(k \ge 1\) be integers, and let \(W_n = C_n \vee K_1\) be the wheel with \(n\) spokes, obtained from the cycle \(C_n\) on \(n\) vertices by adjoining a single hub vertex joined to every vertex of the cycle. Writing \(\pi _k(W_n)\) for the number of distinct proper \(k\)-colourings of \(W_n\), one has
c(v): the number of distinct colours represented at v.
Given a \(k\)-edge colouring \(\mathscr {C}\) of \(G\) we shall denote by \(c(v)\) the number of distinct colours represented at \(v\). Clearly, we always have
Moreover, \(\mathscr {C}\) is a proper \(k\)-edge colouring if and only if equality holds in (6.3) for all vertices \(v\) of \(G\).
Let \(G\) be a finite loopless graph, and let \(\mathscr {C}\) be a \(k\)-edge colouring of \(G\), that is, an arbitrary assignment of one of the colours \(1, 2, \ldots , k\) to each edge (not necessarily proper). For a vertex \(v\), write \(c(v)\) for the number of distinct colours represented at \(v\), and \(d(v)\) for its degree. Then
Component count \(\omega \)(G).
\(\omega \)(G) is the number of components of G — the classes into which the equivalence relation "is connected to" partitions the vertex set. Defined in chapter 1, not chapter 4, so there is no chapter-4 text to quote.
The number of spanning trees \(\tau \)(G): the number of tree subgraphs T \(\le \) G on the full vertex set.
We denote the number of spanning trees of G by \(\tau \)(G).
Let \(G\) be a finite graph, and write \(\tau (G)\) for its number of spanning trees, that is, the number of tree subgraphs of \(G\) on the full vertex set. If \(e\) is an edge of \(G\), then
where \(G - e\) is the graph obtained by deleting \(e\) and \(G \cdot e\) is the graph obtained by contracting \(e\) (deleting it and identifying its two ends).
Let \(G\) be a connected simple graph on a finite vertex set, and let \(k \ge 1\). Suppose that \(G\) has exactly \(2k\) vertices of odd degree. Then there exist \(k\) trails \(Q_1, \dots , Q_k\) in \(G\) such that every edge of \(G\) lies in exactly one of them; that is, the trails are pairwise edge-disjoint and their edge sets together exhaust \(E(G)\).
Let \(G\) be a bipartite graph on a finite vertex set with bipartition into parts \(X\) and \(Y\) whose union is the whole vertex set, and let \(w\) assign a real weight to each edge. Suppose \(l\) is a feasible vertex labelling for these weights, and let \(G_l\) be the corresponding equality subgraph. If \(G_l\) contains a perfect matching \(M^*\), then \(M^*\), regarded as a matching of \(G\), is an optimal matching of \(G\) — that is, a perfect matching of maximum total weight.
Let \(G\) be a finite simple graph on \(\nu \ge 3\) vertices that is not Hamiltonian and has exactly \(\binom {\nu -1}{2}+1\) edges. Then \(G\) is isomorphic to \(C_{1,\nu }\), unless \(\nu = 5\), in which case \(G\) is isomorphic to either \(C_{1,5}\) or \(C_{2,5}\).
Let \(G\) be a finite simple graph, let \(C\) be a cycle in \(G\), and let \(e\) and \(e'\) be edges of \(G\) that do not lie on \(C\); write \(B\) and \(B'\) for the bridges of \(C\) containing \(e\) and \(e'\) respectively. If \(B\) and \(B'\) overlap, then either they are skew, or their sets of vertices of attachment to \(C\) coincide and consist of exactly three vertices — that is, \(B\) and \(B'\) are equivalent \(3\)-bridges.
The (4,4)-Ramsey graph = the Paley graph of order 17.
the \((4, 4)\)-Ramsey graph is obtained by regarding the vertices as elements of the field of integers modulo 17, and joining two vertices if their difference is a quadratic residue of 17 (either 1, 2, 4, 8, 9, 13, 15 or 16). It has been conjectured that the \((k, k)\)-Ramsey graphs are always self-complementary (that is, isomorphic to their complements); this is true for \(k = 2\), 3 and 4.
Let \(P_{17}\) denote the Paley graph of order \(17\): its vertices are the elements of the field \(\mathbb {Z}/17\mathbb {Z}\), and two vertices are joined by an edge whenever their difference is a nonzero quadratic residue modulo \(17\) (that is, one of \(1, 2, 4, 8, 9, 13, 15, 16\)). Then \(P_{17}\) contains no clique on four vertices; equivalently, there is no set of four pairwise adjacent vertices.
Let \(G\) be the Paley graph of order \(17\): its vertices are the elements of the field \(\mathbb {Z}/17\mathbb {Z}\), and two vertices are joined by an edge exactly when their difference is a nonzero quadratic residue modulo \(17\) (that is, one of \(1, 2, 4, 8, 9, 13, 15, 16\)). Then \(G\) contains no independent set of size \(4\); equivalently, among any four vertices of \(G\) some two are adjacent.
Let \(G\) be a finite simple graph on \(n\) vertices with \(m\) edges and minimum degree \(\delta \). Suppose that \(n\) is even, that \(2\delta \lt n\), and that
Then \(G\) has a perfect matching.
Let \(G\) be a finite simple graph that is \(k\)-regular for some integer \(k \ge 1\), and whose edge connectivity satisfies \(\kappa '(G) \ge k - 1\). If \(G\) has an even number of vertices, then \(G\) has a perfect matching, that is, a set of pairwise disjoint edges that together cover every vertex.
Let \(G\) be a finite simple graph that is \(3\)-regular (every vertex has exactly three neighbours) and bridgeless (no edge of \(G\) is a bridge, i.e. removing any single edge leaves the graph with the same number of connected components). Then \(G\) has a perfect matching: there is a set of edges, no two sharing a vertex, that together cover every vertex of \(G\).
For positive integers \(k\) and \(l\), write \(r(a,b)\) for the Ramsey number: the least \(n\) such that every simple graph on \(n\) vertices contains either a clique on \(a\) vertices or an independent set of \(b\) vertices. Then, provided \(k \ge 1\) and \(l \ge 1\),
For integers \(k \ge 2\) and \(l \ge 2\), the Ramsey numbers satisfy
where \(r(k, l)\) denotes the least number of vertices \(n\) such that every simple graph on \(n\) vertices contains either a clique on \(k\) vertices or an independent set on \(l\) vertices.
For natural numbers \(k, l \ge 2\), let \(r(k, l)\) denote the Ramsey number: the least \(n\) such that every simple graph on \(n\) vertices contains either a clique on \(k\) vertices or an independent set on \(l\) vertices. If both \(r(k, l-1)\) and \(r(k-1, l)\) are even, then
For an integer \(k \ge 1\), let \(r(k,k)\) denote the Ramsey number: the least \(n\) such that every simple graph on \(n\) vertices contains either a clique of size \(k\) or an independent set of size \(k\). Then
For any natural numbers \(k_1, \ldots , k_m\), the multicolour Ramsey number satisfies
the right-hand side being the multinomial coefficient \(\binom {k_1 + \cdots + k_m}{k_1, \ldots , k_m}\).
Fix an integer \(m \ge 1\) of colours and integers \(k_1, k_2, \dots , k_m\) with \(k_i \ge 2\) for every \(i\), and let \(r(k_1, \dots , k_m)\) denote the multicolour Ramsey number: the least \(n\) such that every colouring of the edges of the complete graph \(K_n\) with \(m\) colours contains, for some colour \(i\), a complete subgraph on \(k_i\) vertices all of whose edges have colour \(i\). Then
where the \(i\)th summand is obtained from \(r(k_1, \dots , k_m)\) by decreasing the argument \(k_i\) by one.
For natural numbers \(k\) and \(l\), let \(r(k,l)\) denote the Ramsey number: the least \(n\) such that every simple graph on \(n\) vertices contains either a clique on \(k\) vertices or an independent set on \(l\) vertices. Then \(r(k,l) = r(l,k)\).
For integers \(k, l \ge 1\), the Ramsey number \(r(k, l)\) — the least \(n\) such that every simple graph on \(n\) vertices contains either a clique on \(k\) vertices or an independent set on \(l\) vertices — satisfies
Let the Ramsey number \(r(k,l)\) denote the least natural number \(n\) such that every simple graph on \(n\) vertices contains either a clique on \(k\) vertices or an independent set on \(l\) vertices. Then for every integer \(l \ge 1\) one has \(r(1,l) = 1\).
For every integer \(k \ge 1\), the Ramsey number satisfies \(r(k, 1) = 1\); that is, \(1\) is the least number of vertices \(n\) for which every simple graph on \(n\) vertices contains either a clique of size \(k\) or an independent set of size \(1\).
For every integer \(l \ge 1\), the Ramsey number satisfies \(r(2, l) = l\). Equivalently, \(l\) is the least number of vertices \(n\) such that every simple graph on \(n\) vertices contains either an edge (a clique on \(2\) vertices) or an independent set of size \(l\).
The multicolour Ramsey number r(k\(_{1}\),…,k_m).
The least n for which every m-edge colouring of K\(_{n}\) yields a monochromatic complete subgraph on k\(_{i}\) vertices in some colour i — see the verbatim text on IsRamseyBoundMulti above.
For every integer \(n \ge 2\), the \(n\)-colour triangle Ramsey number \(r_n = r(3, 3, \dots , 3)\) — the least \(N\) for which every colouring of the edges of the complete graph on \(N\) vertices with \(n\) colours yields a monochromatic triangle in some colour — satisfies
where \(e = \exp (1)\) and \(\lfloor \, \cdot \, \rfloor \) is the floor function.
For an integer \(n \ge 1\), write \(r_n\) for the \(n\)-colour triangle Ramsey number: the least \(N\) such that every colouring of the edges of the complete graph \(K_N\) with \(n\) colours contains a triangle whose three edges all receive the same colour. Then
Let \(r_3 = r(3,3,3)\) denote the three-colour Ramsey number for triangles, that is, the least \(n\) such that every colouring of the edges of the complete graph \(K_n\) with three colours contains a monochromatic triangle. Then \(r_3 \le 17\); equivalently, every \(3\)-edge-colouring of \(K_{17}\) admits three vertices whose three connecting edges all receive the same colour.
Let \(G\) be a finite simple graph on a vertex set \(V\), let \(s \subseteq V\), and let \(u, v \in s\). If there is a walk \(p\) from \(u\) to \(v\) in \(G\) every vertex of which lies in \(s\), then \(u\) and \(v\), regarded as vertices of the induced subgraph \(G[s]\), are reachable from one another in \(G[s]\); that is, some walk joins them within \(G[s]\).
Let \(G\) be a finite simple graph, and let \(k \ge 1\) be a positive integer such that every vertex of \(G\) has degree exactly \(k\). Suppose that \(G\) is bipartite with parts \(s\) and \(t\); that is, \(s\) and \(t\) are disjoint sets of vertices whose union contains every vertex, and every edge of \(G\) joins a vertex of \(s\) to a vertex of \(t\). Then the two parts have the same size, \(\left\lvert s\right\rvert = \left\lvert t\right\rvert \).
Let \(G\) be a finite simple graph, and suppose \(G\) is regular of degree \(2k\) for some natural number \(k\), meaning every vertex has exactly \(2k\) neighbours. Then \(G\) is \(2\)-factorable: its edges decompose into finitely many pairwise edge-disjoint \(2\)-factors, that is, edge-disjoint \(2\)-regular spanning subgraphs whose union is all of \(G\).
Let \(n \ge 1\), and let \(r_n = r(3,\ldots ,3)\) be the \(n\)-colour Ramsey number for triangles. If \(S_1, S_2, \ldots , S_n\) partition the set \(\{ 1, 2, \ldots , r_n\} \) — that is, each \(S_i\) is a subset of \(\{ 1, \ldots , r_n\} \) and every integer \(k\) with \(1 \le k \le r_n\) lies in exactly one of the \(S_i\) — then for some \(i\) the part \(S_i\) contains integers \(x, y, z\) (not necessarily distinct) with \(x + y = z\).
The Schur number s\(_{n}\). Nonemptiness of the sInf set is discharged by Thm 7.10 (s\(_{n}\) \(\le \) r\(_{n}\)), itself deferred.
Let \(s_n\) denote the least integer such that, in any partition of \(\{ 1,2,\ldots ,s_n\} \) into \(n\) subsets, there is a subset which contains a solution to (7.19). [...] Also, from theorem 7.10 and exercise 7.2.3 we have the upper bound
For each integer \(n \ge 1\), the Schur numbers satisfy
where \(s_m\) denotes the least \(N\) such that every partition of \(\{ 1, 2, \ldots , N\} \) into \(m\) subsets has some part containing elements \(x, y, z\) with \(x + y = z\). Equivalently, \(s_n \ge 3 s_{n-1} - 1\).
The Schur number for three parts equals fourteen, \(s_3 = 14\). Equivalently, \(14\) is the least integer \(N\) with the property that every partition of \(\{ 1,2,\ldots ,N\} \) into three subsets has some subset containing elements \(x,y,z\) (not necessarily distinct) with \(x+y=z\).
Let \(G\) be a self-complementary simple graph on a finite vertex set \(V\); that is, \(G\) is isomorphic to its complement \(\overline{G}\). Then the number of vertices \(\left\lvert V\right\rvert \) satisfies \(\left\lvert V\right\rvert \equiv 0 \pmod4\) or \(\left\lvert V\right\rvert \equiv 1 \pmod4\).
Let \(r(3,3)\) denote the Ramsey number for the pair \((3,3)\): the least natural number \(n\) such that every simple graph on \(n\) vertices contains either three mutually adjacent vertices (a triangle) or three mutually non-adjacent vertices (an independent set of size \(3\)). Then
Two bridges of a cycle are skew: they have attachments a, a’, b, b’ occurring in that cyclic order on c (a, b of one bridge, a’, b’ of the other).
Two bridges B and B’ are skew if there are four distinct vertices u, v, u’ and v’ of C such that u and v are vertices of attachment of B, u’ and v’ are vertices of attachment of B’, and the four vertices appear in the cyclic order u, u’, v, v’ on C.
Let \(G\) be a finite simple graph, and let \(T\) be a spanning tree of \(G\): a connected, acyclic subgraph of \(G\) that touches every vertex. If \(e\) is an edge of \(G\) that does not lie in \(T\), then the graph \(T + e\) obtained by adjoining \(e\) to \(T\) contains a cycle, i.e. there is a vertex \(u\) and a closed walk from \(u\) to itself in \(T + e\) that is a cycle.
Edge subdivision: replace e = uv by a length-2 path through a new vertex.
It is convenient, now, to introduce the operation of subdivision of an edge. An edge \(e\) is said to be subdivided when it is deleted and replaced by a path of length two connecting its ends, the internal vertex of this path being a new vertex.
A k-factor: a k-regular spanning subgraph.
5.1.5 A \(k\)-factor of \(G\) is a \(k\)-regular spanning subgraph of \(G\), and \(G\) is \(k\)-factorable if there are edge-disjoint \(k\)-factors \(H_1, H_2, \ldots , H_n\) such that \(G = H_1 \cup H_2 \cup \ldots \cup H_n\).
A maximum matching: a matching no larger than which exists.
A subset \(M\) of \(E\) is called a matching in \(G\) if its elements are links and no two are adjacent in \(G\); the two ends of an edge in \(M\) are said to be matched under \(M\). A matching \(M\) saturates a vertex \(v\), and \(v\) is said to be \(M\)-saturated, if some edge of \(M\) is incident with \(v\); otherwise, \(v\) is \(M\)-unsaturated. If every vertex of \(G\) is \(M\)-saturated, the matching \(M\) is perfect. \(M\) is a maximum matching if \(G\) has no matching \(M'\) with \(|M'| \gt |M|\); clearly, every perfect matching is maximum.
Type 1: every (k-1)-colouring assigns u and v the same colour.
We shall say that a {u, v}-component G\(_{i}\) of G is of type 1 if every (k-1)-colouring of G\(_{i}\) assigns the same colour to u and v …
Let \(G\) be a simple graph on a finite vertex set, and suppose \(G\) is bipartite, meaning its vertices can be properly colored with two colors so that adjacent vertices receive different colors. Let \(H\) be any subgraph of \(G\) — a simple graph on the same vertex set every edge of which is also an edge of \(G\) — and let \(s\) be any set of vertices. Then the subgraph of \(H\) induced on \(s\), obtained by keeping only the vertices in \(s\) together with those edges of \(H\) joining two vertices of \(s\), is again bipartite.
Let \(G\) be a finite connected graph, and let \(C\) be a longest cycle of \(G\), with its vertices listed in cyclic order as \(v_0, v_1, v_2, \dots \) (so \(v_k\) is the vertex reached after \(k\) steps along \(C\), and \(v_{k+1}\) is its successor). Let \(e\) be an edge of \(G\) that does not lie on \(C\), and let \(B\) be the bridge of \(C\) that contains \(e\); suppose \(B\) has at least one vertex not lying on \(C\). If \(v_i\) and \(v_j\) are two distinct vertices of attachment of \(B\) to \(C\), then \(v_{i+1}\) and \(v_{j+1}\) are not adjacent in \(G\).
Let \(G\) be a simple graph on a finite vertex set \(V\), and for each vertex \(v\) write \(d(v)\) for its degree, the number of vertices adjacent to \(v\). Then
where \(\left\lvert E(G)\right\rvert \) is the number of edges of \(G\).
Let \(G\) be a finite simple graph. For an induced subgraph \(H\) of \(G\), write \(\delta (H)\) for its minimum degree (the least number of neighbours of any vertex of \(H\)). Then the chromatic number of \(G\) satisfies
where the maximum is taken over all induced subgraphs \(H\) of \(G\).
Let \(V\) be a finite vertex set with \(\left\lvert V\right\rvert \ge 3\). Then the complete graph on \(V\) — the simple graph in which every pair of distinct vertices is joined by an edge — is Hamiltonian; that is, it contains a cycle passing through every vertex of \(V\) exactly once.
Let \(G\) be a finite simple graph that is a tree, meaning it is connected and contains no cycles. Then the number of edges of \(G\) is one less than the number of vertices; that is, if \(G\) has \(n\) vertices, then it has exactly \(n - 1\) edges.
Let \(G\) be a finite simple graph and let \(T\) be a tree on \(m\) vertices. If every vertex of \(G\) has degree at least \(m-1\), so that the minimum degree satisfies \(\delta (G) \ge m-1\), then \(G\) contains a copy of \(T\); that is, there is an injective map from the vertices of \(T\) into those of \(G\) that sends every edge of \(T\) to an edge of \(G\), exhibiting \(T\) as a subgraph of \(G\).
Let \(G\) be a tree, that is, a finite connected acyclic graph, and let \(v\) be a vertex of \(G\). Then \(v\) is a cut vertex of \(G\) — meaning \(G\) is connected and deleting \(v\) (together with its incident edges) leaves a disconnected graph — if and only if the degree of \(v\) satisfies \(\deg (v) \gt 1\).
Let \(G\) be a tree on a finite vertex set, and let \(u\) and \(v\) be any two of its vertices. Then there is exactly one path in \(G\) from \(u\) to \(v\); that is, among all walks from \(u\) to \(v\), precisely one has no repeated vertices.
Let \(G\) be a finite simple graph that is not bipartite, with \(n\) vertices and \(m\) edges. If
equivalently \(m \gt \tfrac {(n-1)^2}{4} + 1\), then \(G\) contains a triangle, i.e. three pairwise adjacent vertices. (Erdős)
Let \(V\) be a finite vertex set, let \(m\) be a natural number, and let \(G\) be a simple graph on \(V\) that contains no clique on \(m+1\) vertices. Then the number of edges of \(G\) is at most the number of edges of the Turán graph \(T(\left\lvert V\right\rvert , m)\), the complete \(m\)-partite graph on \(\left\lvert V\right\rvert \) vertices whose parts differ in size by at most one.
Let \(V\) be a finite vertex set, let \(m \geq 1\), and let \(G\) be a simple graph on \(V\) that contains no clique on \(m+1\) vertices. Write \(T\) for the Turán graph \(T(|V|, m)\), the complete \(m\)-partite graph on \(|V|\) vertices whose parts differ in size by at most one. If \(G\) has exactly as many edges as \(T\), then \(G\) is isomorphic to \(T\).
Let \(G\) be a simple graph on a finite vertex set \(V\). For a set \(S \subseteq V\), write \(G - S\) for the graph obtained by deleting the vertices of \(S\) (together with their incident edges), and let \(o(G - S)\) denote the number of odd components of \(G - S\), that is, the connected components with an odd number of vertices. Then \(G\) has a perfect matching if and only if
Let \(G\) be a finite simple graph on a vertex set of at least three vertices. Then the vertex connectivity satisfies \(\kappa (G) \ge 2\) if and only if for every pair of distinct vertices \(u, v\) there exist two internally disjoint \(u\)–\(v\) paths, that is, two paths from \(u\) to \(v\) whose only common vertices are the endpoints \(u\) and \(v\).
Let \(G\) be a finite simple graph whose vertex connectivity satisfies \(\kappa (G) \ge 2\). Then for any two distinct vertices \(u\) and \(v\) of \(G\) there is a cycle of \(G\) whose support contains both \(u\) and \(v\).
Let \(G\) be a finite simple graph with edge connectivity \(\kappa '(G)\). Then \(\kappa '(G) \ge 2\) if and only if for every pair of distinct vertices \(u\) and \(v\) there exist two paths from \(u\) to \(v\) whose edge sets are disjoint.
Let \(G\) be a connected finite simple graph with \(\varepsilon \) edges, and let \(\beta (G)\) denote its covering number, the least number of vertices in a covering of \(G\). Then
A {u,v}-component: G[C \(\cup \) {u,v}] for C a component of G \(-\) {u,v}.
Let S be a vertex cut of a connected graph G, and let the components of G - S have vertex sets V\(_{1}\), V\(_{2}\), …, V_n. Then the subgraphs G\(_{i}\) = G[V\(_{i}\) \(\cup \) S] are called the S-components of G (see figure 8.3). We say that colourings of G\(_{1}\), G\(_{2}\), …, G_n agree on S if, for every v \(\in \) S, vertex v is assigned the same colour in each of the colourings.
(The source prints G\(_{i}\) = G[V \(\cup \) S]; V there is an OCR slip for V\(_{i}\), as the surrounding text and figure 8.3 make clear.)
Vertex connectivity \(\kappa \)(G). If a vertex cut exists it is the minimum cut size; otherwise (complete graphs) it is \(\nu \) \(-\) 1.
If \(G\) has at least one pair of distinct nonadjacent vertices, the connectivity \(\kappa (G)\) of \(G\) is the minimum \(k\) for which \(G\) has a \(k\)-vertex cut; otherwise, we define \(\kappa (G)\) to be \(\nu - 1\). Thus \(\kappa (G) = 0\) if \(G\) is either trivial or disconnected.
Let \(G\) be a finite simple graph on a vertex set \(V\), and let \(e\) be an unordered pair of vertices; write \(G - e\) for the graph obtained from \(G\) by deleting the edge \(e\) (leaving \(G\) unchanged if \(e\) is not an edge of \(G\)). Then the vertex connectivities satisfy
Let \(G\) be a simple graph on a finite vertex set, with maximum degree \(\Delta (G)\), and let \(\chi '(G)\) denote its chromatic index — that is, the chromatic number of the line graph of \(G\), equivalently the least number of colours needed to colour the edges of \(G\) so that any two edges sharing an endpoint receive distinct colours. Then either \(\chi '(G) = \Delta (G)\) or \(\chi '(G) = \Delta (G) + 1\).
Let \(G\) be a finite simple graph with maximum degree \(\Delta (G)\), and let \(L(G)\) denote its line graph. Then the chromatic number of \(L(G)\) — equivalently, the chromatic index (edge chromatic number) \(\chi '(G)\) of \(G\) — satisfies
M-alternating walk: consecutive edges alternate in/out of M.
Let \(M\) be a matching in \(G\). An \(M\)-alternating path in \(G\) is a path whose edges are alternately in \(E \backslash M\) and \(M\). For example, the path \(v_5 v_8 v_1 v_7 v_6\) in the graph of figure 5.1\(a\) is an \(M\)-alternating path.
Let \(G\) be a finite simple graph on \(n\) vertices, and enumerate its vertices as \(v_1, v_2, \dots , v_n\) so that their degrees are arranged in non-increasing order, \(d(v_1) \ge d(v_2) \ge \dots \ge d(v_n)\). Then the chromatic number of \(G\) satisfies
The wheel with n spokes: C_n \(\lor \) K\(_{1}\).
(B&M exercise 2.4.2, verbatim; used in exercise 8.4.5(b)). A wheel is a graph obtained from a cycle by adding a new vertex and edges joining it to all the vertices of the cycle; the new edges are called the spokes of the wheel.
For every finite nonempty simple graph \(G\), the vertex connectivity, edge connectivity, and minimum degree satisfy
where \(\kappa (G)\) is the least size of a vertex cut (or \(|V(G)| - 1\) when \(G\) has no vertex cut), \(\kappa '(G)\) is the least size of an edge cut, and \(\delta (G)\) is the minimum degree of a vertex of \(G\).
Let \(G\) be a loopless multigraph with vertex set \(V(G)\), and let \(v \in V(G)\) be a vertex whose deletion leaves at least one other vertex, i.e. \(V(G) \setminus \{ v\} \) is nonempty. Then the singleton \(S = \{ v\} \) is a cut of \(G\), and the size of this cut — the number of edges with exactly one endpoint in \(\{ v\} \), counted with multiplicity — equals the degree \(\deg (v)\) of \(v\).
Fix \(\tau \) with \(0 \lt \tau \lt 1/2\), and let \((t_n)\) be a sequence of natural numbers with \(t_n/n \to \tau \). Then there is a real sequence \((a_n)\) with the following three properties. First, for all sufficiently large \(n\), the number \(a_n\) lies strictly below every real zero of the Krawtchouk polynomial \(K_{t_n}^{(n)}\); that is, \(K_{t_n}^{(n)}(\xi ) = 0\) implies \(a_n \lt \xi \). Second, \(a_n/n \to \tfrac 12 - \sqrt{\tau (1-\tau )}\). Third, for all sufficiently large \(n\), the truncated Christoffel–Darboux kernel satisfies \(\Phi _{n,t_n}(a_n, x) \le 0\) for every integer \(x\) with \(\big\lfloor (\tfrac 12 - \sqrt{\tau (1-\tau )})\, n \big\rfloor \le x \le n\).
Let \(\mathrm{val}_{f_{L,R}}\), \(N\), \(D\), and \(\lVert A\rVert \) be real numbers with \(D \gt 0\) and \(N \gt 0\). If
then \(\mathrm{val}_{f_{L,R}} \le \frac{N}{D}\, \lVert A\rVert \).
For a bit string \(x\), its decoding \(\mathrm{dec}(x)\) is the CNF formula returned by the exact-consumption parser on \(x\) when that parse succeeds, and the empty formula (no clauses) otherwise. Here the parser reads \(x\), with recursion fuel \(|x|\), according to the grammar \(\langle \text{formula}\rangle ::= 0 \mid 1\, \langle \text{clause}\rangle \, \langle \text{formula}\rangle \), \(\langle \text{clause}\rangle ::= 0 \mid \langle \text{literal}\rangle \, \langle \text{clause}\rangle \), \(\langle \text{literal}\rangle ::= 1^{k+1}\, 0\, b\) (the literal with variable index \(k \in \mathbb {N}\) and polarity \(b \in \{ 0,1\} \)), and succeeds only if it consumes all of \(x\).
For every CNF formula \(\varphi \) over variables indexed by \(\mathbb {N}\), \(\mathrm{dec}(\mathrm{enc}(\varphi )) = \varphi \). Here a literal \((v,b)\) is encoded as \(1^{v+1}0b\), a clause as the concatenation of its literal encodings followed by \(0\), and a formula \((C_1,\dots ,C_k)\) as \(1\, \mathrm{enc}(C_1)\cdots 1\, \mathrm{enc}(C_k)\, 0\); and \(\mathrm{dec}(x)\) is the formula obtained by parsing the bit string \(x\), with recursion fuel \(|x|\), according to the grammar \(\langle \text{formula}\rangle ::= 0 \mid 1\, \langle \text{clause}\rangle \, \langle \text{formula}\rangle \), \(\langle \text{clause}\rangle ::= 0 \mid \langle \text{literal}\rangle \, \langle \text{clause}\rangle \), \(\langle \text{literal}\rangle ::= 1^{k+1}\, 0\, b\) (the literal \((k,b)\)) when this parse consumes all of \(x\), and the empty formula otherwise.
Let \(\varphi = (C_1, \dots , C_k)\) be a CNF formula over variables indexed by \(\mathbb {N}\), each clause a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). Its De Morgan dual is the DNF formula \((C_1', \dots , C_k')\), where the term \(C_i'\) is obtained from \(C_i\) by replacing every literal \((v,b)\) with \((v, \neg b)\), keeping the order. Thus if \(\varphi = \bigwedge _i \bigvee _j \ell _{ij}\), its dual is the DNF formula \(\bigvee _i \bigwedge _j \neg \ell _{ij}\); no distributive expansion takes place, and the dual has the same size as \(\varphi \).
Let \(\varphi \) be a CNF formula over variables indexed by \(\mathbb {N}\), a finite list of clauses each a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). Forming its De Morgan dual DNF formula, by replacing every literal \((v,b)\) with \((v,\neg b)\) and reading the clauses as terms, and then forming the De Morgan dual CNF formula of that DNF formula, by again replacing every literal \((v,b)\) with \((v,\neg b)\) and reading the terms as clauses, returns \(\varphi \).
Let \(\varphi = (C_1, \dots , C_k)\) be a CNF formula over variables indexed by \(\mathbb {N}\), each clause a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \), and let \(\varphi ^{\ast }\) be the DNF formula whose terms are obtained from \(C_1, \dots , C_k\) by replacing every literal \((v,b)\) with \((v, \neg b)\). Then for every assignment \(a : \mathbb {N} \to \{ 0,1\} \),
where \(\varphi (a)\) is true when every clause of \(\varphi \) contains a literal \((v,b)\) with \(a(v) = b\), and \(\varphi ^{\ast }(a)\) is true when some term of \(\varphi ^{\ast }\) has all of its literals \((v,b)\) satisfying \(a(v) = b\).
The fallback CNF formula over variables indexed by \(\mathbb {N}\) is the empty formula, consisting of no clauses; it is therefore satisfied by every assignment. It is the fixed formula to which bit strings that fail to parse are decoded, following [AB09, footnote 3].
Let \(\varphi \) be a CNF formula whose variables are indexed by natural numbers. Its variable count \(N(\varphi )\) is one plus the largest index of a variable occurring in some literal of \(\varphi \), and \(0\) if no literal occurs in \(\varphi \) (for instance when \(\varphi \) has no clauses, or only empty clauses). Equivalently, \(N(\varphi )\) is the least \(N \in \mathbb {N}\) such that every variable occurring in \(\varphi \) has index less than \(N\).
For every bit string \(x\), the decoded formula satisfies \(N(\mathrm{dec}(x)) \le |x|\). Here the variable count \(N(\varphi )\) of a CNF formula \(\varphi \) is one plus the largest variable index occurring in a literal of \(\varphi \), or \(0\) if no literal occurs, and \(\mathrm{dec}(x)\) is the formula obtained by parsing \(x\), with recursion fuel \(|x|\), according to the grammar \(\langle \text{formula}\rangle ::= 0 \mid 1\, \langle \text{clause}\rangle \, \langle \text{formula}\rangle \), \(\langle \text{clause}\rangle ::= 0 \mid \langle \text{literal}\rangle \, \langle \text{clause}\rangle \), \(\langle \text{literal}\rangle ::= 1^{k+1}\, 0\, b\) (the literal with variable index \(k \in \mathbb {N}\) and polarity \(b \in \{ 0,1\} \)) when this parse consumes all of \(x\), and the empty formula otherwise.
Let \(\varphi \) be a CNF formula over variables indexed by \(\mathbb {N}\) and let \(n \in \mathbb {N}\). If every literal \((v,b)\) of every clause of \(\varphi \) satisfies \(v + 1 \le n\), then \(N(\varphi ) \le n\), where the variable count \(N(\varphi )\) is one plus the largest variable index occurring in a literal of \(\varphi \), or \(0\) if \(\varphi \) contains no literal.
For a bit string \(x\), the parser runs the fuel-bounded formula parser with fuel \(|x|\) on \(x\) and returns the resulting CNF formula \(\varphi \) if that run succeeds with empty remainder, i.e. consumes all of \(x\); it fails otherwise, in particular whenever trailing bits remain. The formula parser reads the grammar \(\langle \text{formula}\rangle ::= 0 \mid 1\, \langle \text{clause}\rangle \, \langle \text{formula}\rangle \), \(\langle \text{clause}\rangle ::= 0 \mid \langle \text{literal}\rangle \, \langle \text{clause}\rangle \), \(\langle \text{literal}\rangle ::= 1^{k+1}\, 0\, b\) (read as the literal with variable index \(k \in \mathbb {N}\) and polarity \(b \in \{ 0,1\} \)), where the fuel bounds the recursion: with fuel \(g+1\), after reading the leading \(1\) of a clause the formula parser parses that clause and the remaining formula each with fuel \(g\), and after reading a literal the clause parser parses the rest of the clause with fuel \(g\); with fuel \(0\) only a terminating \(0\) can be read.
For every CNF formula \(\varphi \) over variables indexed by \(\mathbb {N}\), the exact-consumption parser applied to the binary encoding \(\mathrm{enc}(\varphi )\) succeeds and returns \(\varphi \). Here a literal \((v,b)\) is encoded as \(1^{v+1}0b\), a clause as the concatenation of its literal encodings followed by \(0\), and a formula \((C_1,\dots ,C_k)\) as \(1\, \mathrm{enc}(C_1)\cdots 1\, \mathrm{enc}(C_k)\, 0\); the parser on a bit string \(x\) reads \(x\), with recursion fuel \(|x|\), according to the grammar \(\langle \text{formula}\rangle ::= 0 \mid 1\, \langle \text{clause}\rangle \, \langle \text{formula}\rangle \), \(\langle \text{clause}\rangle ::= 0 \mid \langle \text{literal}\rangle \, \langle \text{clause}\rangle \), \(\langle \text{literal}\rangle ::= 1^{k+1}\, 0\, b\) (the literal \((k,b)\)), and succeeds only if it consumes all of \(x\).
For a fuel bound \(f \in \mathbb {N}\) and a bit string \(x\), the clause parser with fuel \(f\) is defined by recursion on \(f\). If \(x = 0y\), it succeeds with the empty clause and remainder \(y\). If \(f = g + 1\) and \(x\) begins with \(1\), it reads one literal from the front of \(x\), which succeeds only when \(x = 1^{k+1}0by\) and then gives the literal \((k,b)\) and remainder \(y\); it then runs the clause parser with fuel \(g\) on \(y\), and if that returns a clause \(C\) with remainder \(z\), the result is the clause obtained by prepending \((k,b)\) to \(C\), with remainder \(z\). In every other case (empty input, fuel \(0\) on an input beginning with \(1\), or a failed sub-parse) it fails.
Let \(f \in \mathbb {N}\), let \(x\) and \(r\) be bit strings and let \(C\) be a clause of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). Suppose the clause parser with fuel \(f\) succeeds on \(x\) and returns \(C\) with remainder \(r\). Then \(|r| \le |x|\), and every literal \((v,b)\) of \(C\) satisfies \(v + 1 + |r| \le |x|\). Here the clause parser with fuel \(f\) on \(x\) returns the empty clause and remainder \(y\) when \(x = 0y\); when \(f = g+1\) and \(x = 1^{k+1}0cy\) it prepends the literal \((k,c)\) to the clause returned by the clause parser with fuel \(g\) on \(y\), keeping that call’s remainder and failing if it fails; and it fails on every other input.
Let \(C = (\ell _1, \dots , \ell _m)\) be a clause of literals \(\ell _i = (v_i, b_i)\) with \(v_i \in \mathbb {N}\) and \(b_i \in \{ 0,1\} \), and let \(\mathrm{enc}(C) = 1^{v_1+1}0b_1\cdots 1^{v_m+1}0b_m\, 0\) be its binary encoding. For every bit string \(r\) and every fuel \(f \ge |\mathrm{enc}(C)|\), the clause parser with fuel \(f\) succeeds on \(\mathrm{enc}(C)\, r\) and returns \(C\) with remainder \(r\). Here the clause parser with fuel \(f\) on an input \(x\) returns the empty clause and remainder \(y\) when \(x = 0y\); when \(f = g+1\) and \(x = 1^{k+1}0cy\) it prepends the literal \((k,c)\) to the clause returned by the clause parser with fuel \(g\) on \(y\), keeping that call’s remainder and failing if it fails; and it fails on every other input.
For a fuel bound \(f \in \mathbb {N}\) and a bit string \(x\), the formula parser with fuel \(f\) is defined by recursion on \(f\). If \(x = 0y\), it succeeds with the empty formula (no clauses) and remainder \(y\). If \(f = g+1\) and \(x = 1y\), it runs the clause parser with fuel \(g\) on \(y\) and, if this returns a clause \(C\) with remainder \(z\), runs the formula parser with fuel \(g\) on \(z\); if that returns a formula \(\psi \) with remainder \(w\), the result is the formula whose first clause is \(C\) followed by the clauses of \(\psi \), with remainder \(w\). In all other cases it fails. Here the clause parser with fuel \(g\) on a bit string \(u\) returns the empty clause and remainder \(u'\) when \(u = 0u'\); when \(g = h+1\) and \(u = 1^{j+1}0cu'\) it prepends the literal \((j,c)\) to the clause returned by the clause parser with fuel \(h\) on \(u'\), keeping that call’s remainder; and it fails otherwise.
Let \(f \in \mathbb {N}\), let \(x\) and \(r\) be bit strings and let \(\varphi \) be a CNF formula over variables indexed by \(\mathbb {N}\). Suppose the formula parser with fuel \(f\) succeeds on \(x\) and returns \(\varphi \) with remainder \(r\). Then \(|r| \le |x|\), and every literal \((v,b)\) of every clause of \(\varphi \) satisfies \(v + 1 + |r| \le |x|\). Here the formula parser with fuel \(f\) on \(x\) returns the empty formula and remainder \(y\) when \(x = 0y\); when \(f = g+1\) and \(x = 1y\) it parses a clause \(C\) from \(y\) with fuel \(g\), leaving \(z\), then a formula \(\psi \) from \(z\) with fuel \(g\), leaving \(w\), and returns \(C\) followed by the clauses of \(\psi \) with remainder \(w\); and it fails otherwise. The clause parser with fuel \(g\) on \(u\) returns the empty clause and remainder \(u'\) when \(u = 0u'\), and when \(g = h+1\) and \(u = 1^{j+1}0cu'\) it prepends the literal \((j,c)\) to the clause parsed from \(u'\) with fuel \(h\), keeping that remainder; it fails otherwise.
Let \(\varphi = (C_1,\dots ,C_k)\) be a CNF formula over variables indexed by \(\mathbb {N}\), and let \(\mathrm{enc}(\varphi ) = 1\, \mathrm{enc}(C_1)\cdots 1\, \mathrm{enc}(C_k)\, 0\) be its binary encoding, where a clause with literals \((v_1,b_1),\dots ,(v_m,b_m)\) is encoded as \(1^{v_1+1}0b_1\cdots 1^{v_m+1}0b_m\, 0\). For every bit string \(r\) and every fuel \(f \ge |\mathrm{enc}(\varphi )|\), the formula parser with fuel \(f\) succeeds on \(\mathrm{enc}(\varphi )\, r\) and returns \(\varphi \) with remainder \(r\). Here the formula parser with fuel \(f\) on an input \(x\) returns the empty formula and remainder \(y\) when \(x = 0y\); when \(f = g+1\) and \(x = 1y\) it parses a clause \(C\) from \(y\) with fuel \(g\), leaving \(z\), then a formula \(\psi \) from \(z\) with fuel \(g\), leaving \(w\), and returns \(C\) followed by the clauses of \(\psi \) with remainder \(w\); and it fails otherwise. The clause parser with fuel \(g\) on \(u\) returns the empty clause and remainder \(u'\) when \(u = 0u'\), and when \(g = h+1\) and \(u = 1^{j+1}0cu'\) it prepends the literal \((j,c)\) to the clause parsed from \(u'\) with fuel \(h\), keeping that remainder; it fails otherwise.
For a bit string \(x\), the literal parser succeeds exactly when \(x = 1^{k+1}\, 0\, b\, y\) for some \(k \in \mathbb {N}\), bit \(b \in \{ 0,1\} \) and bit string \(y\), in which case it returns the literal \((k, b)\), with variable index \(k\) and polarity \(b\), together with the unconsumed remainder \(y\). On every other input (one not beginning with \(1\), or one ending before the \(0\) and the polarity bit that follow its leading run of ones) it fails.
Let \(x\) and \(r\) be bit strings and let \((v, b)\) be a literal with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). If the literal parser, which on an input \(1^{k+1}0cy\) returns the literal \((k,c)\) with remainder \(y\) and fails on every other input, returns \((v,b)\) with remainder \(r\) when applied to \(x\), then \(v + 3 + |r| = |x|\).
Let \(\ell = (v, b)\) be a literal with variable index \(v \in \mathbb {N}\) and polarity bit \(b \in \{ 0,1\} \), whose binary encoding is \(1^{v+1}0b\), and let \(r\) be any bit string. The literal parser, which on an input of the form \(1^{k+1}0cy\) returns the literal \((k,c)\) with remainder \(y\) and fails on every other input, succeeds on \(1^{v+1}0b\, r\) and returns \(\ell \) with remainder \(r\).
Let \(\alpha \) be a type of variables and let \(\varphi \) be a CNF formula over \(\alpha \), that is, a conjunction of clauses, each a disjunction of literals \(x_v^b\) (\(v \in \alpha \), \(b \in \{ 0,1\} \)), where \(x_v^b\) is satisfied by an assignment \(a\) exactly when \(a(v) = b\). The formula \(\varphi \) is satisfiable if there exists an assignment \(a : \alpha \to \{ 0,1\} \) under which \(\varphi \) evaluates to true [AB09, Section 2.3.1].
Let \(\varphi = (C_1, \dots , C_k)\) be a CNF formula over variables indexed by \(\mathbb {N}\), given as a finite list of clauses, each a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). Its binary encoding is
where a clause \((\ell _1,\dots ,\ell _m)\) is encoded as \(\mathrm{enc}(\ell _1)\cdots \mathrm{enc}(\ell _m)\, 0\) and a literal \((v,b)\) as \(1^{v+1}0b\). At the formula level a \(1\) announces another clause and a \(0\) ends the formula; in particular the empty formula is encoded as the single bit \(0\).
Let \(C = (\ell _1, \dots , \ell _m)\) be a clause, i.e. a finite list of literals \(\ell _i = (v_i, b_i)\) with variable indices \(v_i \in \mathbb {N}\) and polarity bits \(b_i \in \{ 0,1\} \). Its binary encoding is the bit string
the concatenation of the literal encodings \(1^{v_i+1}0b_i\) followed by a closing \(0\); the closing bit is unambiguous because every literal encoding begins with \(1\).
For a literal \(\ell = (v, b)\), consisting of a variable index \(v \in \mathbb {N}\) and a polarity bit \(b \in \{ 0,1\} \), its binary encoding is the bit string \(\mathrm{enc}(\ell ) = 1^{v+1}\, 0\, b\) of length \(v+3\) (writing \(1\) for true and \(0\) for false): the index in unary as \(v+1\) ones, nonempty even when \(v = 0\), then a terminating \(0\), then the polarity bit \(b\).
Let \(\varphi = (C_1, \dots , C_k)\) be a CNF formula over variables indexed by \(\mathbb {N}\), each clause a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \), and let \(\varphi ^{\ast }\) be the DNF formula whose terms are obtained from \(C_1, \dots , C_k\) by replacing every literal \((v,b)\) with \((v, \neg b)\). Then \(\varphi ^{\ast }\) is a tautology, i.e. for every assignment \(a : \mathbb {N} \to \{ 0,1\} \) some term of \(\varphi ^{\ast }\) has all of its literals \((v,b)\) satisfying \(a(v) = b\), if and only if \(\varphi \) is unsatisfiable, i.e. no assignment makes every clause of \(\varphi \) contain a literal \((v,b)\) with \(a(v) = b\) [AB09, Section 2.6.1, Example 2.21].
Let \(\alpha \) be a type of variables, let \(\varphi \) be a CNF formula over \(\alpha \), and let \(k \in \mathbb {N}\). The formula \(\varphi \) has width at most \(k\) (is a \(k\)CNF) if every clause of \(\varphi \) contains at most \(k\) literals [AB09, Section 2.3.1]. In particular, a formula with no clauses has width at most \(k\) for every \(k\).
A DNF formula over variables indexed by \(\mathbb {N}\) is a finite list of terms \((T_1, \dots , T_k)\), where each term \(T_i\) is a finite list of literals \((v, b)\) with variable index \(v \in \mathbb {N}\) and polarity bit \(b \in \{ 0,1\} \). Each term is read as the conjunction of its literals and the formula as the disjunction of its terms; the syntax is the same list-of-literal-lists shape as a CNF formula over \(\mathbb {N}\), but it carries this dual reading.
For a bit string \(x\), its DNF decoding is the DNF formula whose list of terms is the list of clauses of the CNF decoding \(\mathrm{dec}(x)\). Here \(\mathrm{dec}(x)\) is the list of literal lists obtained by parsing \(x\), with recursion fuel \(|x|\), according to the grammar \(\langle \text{formula}\rangle ::= 0 \mid 1\, \langle \text{clause}\rangle \, \langle \text{formula}\rangle \), \(\langle \text{clause}\rangle ::= 0 \mid \langle \text{literal}\rangle \, \langle \text{clause}\rangle \), \(\langle \text{literal}\rangle ::= 1^{k+1}\, 0\, b\) (the literal with variable index \(k \in \mathbb {N}\) and polarity \(b \in \{ 0,1\} \)) when this parse consumes all of \(x\), and the empty list otherwise; so a malformed string decodes to the DNF formula with no terms.
Let \(\varphi = (C_1, \dots , C_k)\) be a CNF formula over variables indexed by \(\mathbb {N}\), each clause a list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \), with binary encoding \(\mathrm{enc}(\varphi ) = 1\, \mathrm{enc}(C_1)\cdots 1\, \mathrm{enc}(C_k)\, 0\), where a clause is encoded as the concatenation of its literal encodings \(1^{v+1}0b\) followed by \(0\). Then decoding \(\mathrm{enc}(\varphi )\) as a DNF formula yields the DNF formula whose terms are exactly \(C_1, \dots , C_k\). Here a bit string \(x\) is decoded as the DNF formula whose terms are the literal lists obtained by parsing \(x\), with recursion fuel \(|x|\), according to the grammar \(\langle \text{formula}\rangle ::= 0 \mid 1\, \langle \text{clause}\rangle \, \langle \text{formula}\rangle \), \(\langle \text{clause}\rangle ::= 0 \mid \langle \text{literal}\rangle \, \langle \text{clause}\rangle \), \(\langle \text{literal}\rangle ::= 1^{k+1}\, 0\, b\) when this parse consumes all of \(x\), and as the formula with no terms otherwise.
For every DNF formula \(\psi \) over variables indexed by \(\mathbb {N}\), decoding its binary encoding as a DNF formula returns \(\psi \). Here a DNF formula \((T_1,\dots ,T_k)\), each term a list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \), is encoded as \(1\, \mathrm{enc}(T_1)\cdots 1\, \mathrm{enc}(T_k)\, 0\), a term as the concatenation of its literal encodings \(1^{v+1}0b\) followed by \(0\); and a bit string \(x\) is decoded as the DNF formula whose terms are the literal lists obtained by parsing \(x\), with recursion fuel \(|x|\), according to the grammar \(\langle \text{formula}\rangle ::= 0 \mid 1\, \langle \text{clause}\rangle \, \langle \text{formula}\rangle \), \(\langle \text{clause}\rangle ::= 0 \mid \langle \text{literal}\rangle \, \langle \text{clause}\rangle \), \(\langle \text{literal}\rangle ::= 1^{k+1}\, 0\, b\) (the literal \((k,b)\)) when this parse consumes all of \(x\), and as the formula with no terms otherwise.
Let \(\psi = (T_1, \dots , T_k)\) be a DNF formula over variables indexed by \(\mathbb {N}\), each term a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). Its De Morgan dual is the CNF formula \((T_1', \dots , T_k')\), where the clause \(T_i'\) is obtained from \(T_i\) by replacing every literal \((v,b)\) with \((v, \neg b)\), keeping the order. Thus if \(\psi = \bigvee _i \bigwedge _j \ell _{ij}\), its dual is the CNF formula \(\bigwedge _i \bigvee _j \neg \ell _{ij}\), of the same size.
Let \(\psi \) be a DNF formula over variables indexed by \(\mathbb {N}\), a finite list of terms each a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). Forming its De Morgan dual CNF formula, by replacing every literal \((v,b)\) with \((v,\neg b)\) and reading the terms as clauses, and then forming the De Morgan dual DNF formula of that CNF formula, by again replacing every literal \((v,b)\) with \((v,\neg b)\) and reading the clauses as terms, returns \(\psi \).
Let \(\psi = (T_1, \dots , T_k)\) be a DNF formula over variables indexed by \(\mathbb {N}\), each term \(T_i\) a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \), and let \(a : \mathbb {N} \to \{ 0,1\} \) be an assignment. The value \(\psi (a) \in \{ 0,1\} \) is true exactly when some term \(T_i\) has all of its literals satisfied by \(a\), a literal \((v,b)\) being satisfied when \(a(v) = b\):
In particular the formula with no terms evaluates to false under every assignment, and an empty term is satisfied by every assignment [AB09, Section 2.6.1].
Let \(\psi = (T_1, \dots , T_k)\) be a DNF formula over variables indexed by \(\mathbb {N}\), each term a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). Its binary encoding is the binary encoding of the same list of literal lists read as a CNF formula,
where a term \((\ell _1, \dots , \ell _m)\) is encoded as \(\mathrm{enc}(\ell _1)\cdots \mathrm{enc}(\ell _m)\, 0\) and a literal \((v,b)\) as \(1^{v+1}0b\).
Let \(\psi = (T_1, \dots , T_k)\) be a DNF formula over variables indexed by \(\mathbb {N}\), each term a finite list of literals \((v,b)\) with \(v \in \mathbb {N}\) and \(b \in \{ 0,1\} \). The formula \(\psi \) is a tautology if for every assignment \(a : \mathbb {N} \to \{ 0,1\} \) some term \(T_i\) has all of its literals satisfied by \(a\), where \((v,b)\) is satisfied when \(a(v) = b\); that is, \(\psi \) evaluates to true under every assignment [AB09, Section 2.6.1].
Let \(Z\) be a measurable real random variable on a probability space \((\Omega , \mu )\) which is strictly sub-Gaussian with variance proxy \(\sigma ^2 \gt 0\), i.e. for every real \(t\) the function \(e^{tZ}\) is \(\mu \)-integrable and \(\mathbb {E}_\mu [e^{tZ}] \le e^{\sigma ^2 t^2/2}\), and whose second moment is exactly \(\mathbb {E}_\mu [Z^2] = \sigma ^2\). Then the centered square \(Z^2 - \sigma ^2\) satisfies the Bernstein MGF condition ProbabilityTheory.HasBernsteinMGF with parameters \((2\sigma ^4,\, 1/(4\sigma ^2))\): for every \(t\) with \(\left\lvert t\right\rvert \le 1/(4\sigma ^2)\), \(e^{t(Z^2 - \sigma ^2)}\) is \(\mu \)-integrable and \(\mathbb {E}_\mu [e^{t(Z^2 - \sigma ^2)}] \le e^{2\sigma ^4 t^2}\). This is the sub-exponential MGF bound of Vershynin, Proposition 2.7.1(e), for the square of a sub-Gaussian variable (the sub-Gaussian/sub-exponential link of Lemma 2.7.6), i.e. the scalar case of the Hanson–Wright circle of results (Theorem 6.2.1), with the sharp Gaussian-case constants; the Gaussian decoupling identity it rests on is the one from Vershynin’s Lemma 6.2.2.
The weight function on the augmented universe \(U \oplus \mathrm{Bool}\) built from a Subset Sum instance \((w, T)\). Writing \(W = \sum _{a} w(a)\), it sends the original item \(\mathrm{inl}\, a\) to \(w(a)\), the first dummy item \(\mathrm{inr}\, \mathsf{true}\) to \(2W - T\), and the second dummy item \(\mathrm{inr}\, \mathsf{false}\) to \(W + T\).
Let \(w : U \to \mathbb {N}\) be a weight function on a finite type \(U\), and let \(T \in \mathbb {N}\) satisfy \(T \le \sum _{a} w(a)\). If the Subset Sum instance \((w, T)\) has a solution — some subset of \(U\) with total weight exactly \(T\) — then the Partition instance on the augmented type \(U \oplus \mathrm{Bool}\) whose weights are given by the reduction weight function built from \((w, T)\) has a solution; that is, \(U \oplus \mathrm{Bool}\) admits a subset whose total weight equals that of its complement.
Let \(U\) be a finite type, let \(w : U \to \mathbb {N}\) be a weight function, and let \(T \in \mathbb {N}\) be a target satisfying \(T \le \sum _{a} w(a)\). Write \(W = \sum _{a} w(a)\), and form the weight function on the augmented universe \(U \oplus \mathrm{Bool}\) that assigns weight \(w(a)\) to each original item \(a\), weight \(2W - T\) to one added dummy item, and weight \(W + T\) to a second added dummy item. Then some subset of \(U\) has total \(w\)-weight exactly \(T\) if and only if the augmented universe admits a partition into two parts of equal total weight.
Let \(w : \mathbb {N}^{U}\) be a weight function on a finite type \(U\), and let \(T \in \mathbb {N}\) be a target with \(T \le \sum _{a} w(a)\). Form the reduction weight function on the augmented type \(U \oplus \mathrm{Bool}\), which assigns each original item the weight \(w(a)\), the first dummy item the weight \(2\sum _a w(a) - T\), and the second dummy item the weight \(\sum _a w(a) + T\). Then, if this augmented instance is a yes-instance of Partition — some subset of \(U \oplus \mathrm{Bool}\) has the same total weight as its complement — the original instance \((w, T)\) is a yes-instance of Subset Sum, i.e. some subset of \(U\) has total weight exactly \(T\).
Let \(m\) be a finite multiset with elements drawn from a type \(\alpha \), and let \(s\) be a finite set of elements of \(\alpha \) that contains every element occurring in \(m\). Writing \(\mathrm{count}_m(a)\) for the multiplicity of \(a\) in \(m\), the multiplicities over \(s\) sum to the total size of the multiset:
Let \(G\) be a loopless multigraph, given by a finite vertex set \(V\) together with a multiset \(E\) of edges, each edge an unordered pair of distinct vertices of \(V\). For a vertex \(v\), let \(\deg (v)\) be the number of edges of \(G\) incident to \(v\), counted with multiplicity. Then
where \(\left\lvert E\right\rvert \) is the number of edges counted with multiplicity.
Let \(D\) be a digraph on a finite vertex set \(V\), and write \(\varepsilon \) for its number of arcs. Then both the sum of the indegrees and the sum of the outdegrees over all vertices equal \(\varepsilon \):
Let \((\Omega ,\mu )\) be a probability space, and let \((Y_j)_{j}\) be a family of independent, measurable real-valued random variables, each distributed as a centred Gaussian with mean \(0\) and variance \(\sigma _0 \ge 0\). Then for any real coefficients \((x_j)_j\) and any finite index set \(s\), the random variable \(\sum _{j \in s} x_j\, Y_j\) has distribution \(N\! \Bigl(0,\; \bigl(\sum _{j \in s} x_j^2\bigr)\, \sigma _0\Bigr)\); that is, its law is the centred Gaussian on \(\mathbb {R}\) with variance \(\sigma _0 \sum _{j \in s} x_j^2\).
Fix \(n\), a finset \(S \subseteq \mathrm{Fin}\, n\), and an assignment \(f : S \to \{ X,Y,Z\} \) of a non-identity Pauli to each coordinate in \(S\). Let \(p\) be the Pauli string of length \(n\) that equals \(f(i)\) at each \(i \in S\) and equals the identity \(I\) at every coordinate outside \(S\). Then the support of \(p\) — the set of coordinates where \(p\) differs from \(I\) — is exactly \(S\).
Let \(n\) and \(k\) be natural numbers with \(k \le n\). Among all restrictions on \(n\) variables — maps assigning each variable either a fixed Boolean value or leaving it free — those whose set of free variables has cardinality exactly \(k\) number precisely \(\binom {n}{k}\, 2^{\, n-k}\).
Let \(f\) be a Boolean function on \(n\) variables and fix natural numbers \(d\) and \(k\). For a restriction \(\rho \) on these variables, write \(f|_\rho \) for the restricted function that sends a point \(x\) to the value of \(f\) obtained by fixing each assigned coordinate as \(\rho \) prescribes and reading each free coordinate off \(x\), and call \(\rho \) bad for \(f\) at depth \(d\) when the decision-tree depth of \(f|_\rho \) is strictly greater than \(d\). Then the number of restrictions that have exactly \(k\) free variables and are bad for \(f\) at depth \(d\) is the same whether these are counted within the collection of all restrictions with exactly \(k\) free variables or within the collection of all restrictions on the \(n\) variables that both have exactly \(k\) free variables and are bad for \(f\) at depth \(d\).
Let \(f\) be a CNF formula on \(n \ge 1\) variables whose width is at most \(w\), where \(w \ge 1\), and suppose every clause of \(f\) is proper: within any one clause, two literals sharing the same variable are equal, and no literal is repeated. Let \(p \in \mathbb {R}\) satisfy \(0 \lt p \le \tfrac {1}{40w}\), and in addition \(p \le 1\). Draw a Bernoulli(\(p\)) random restriction \(\rho \), which independently leaves each variable free with probability \(p\) and otherwise fixes it to one of the two Boolean values, each with probability \((1-p)/2\); write \(f\vert _\rho \) for the function of the remaining free variables obtained by fixing the assigned coordinates according to \(\rho \). Then for every \(t \in \mathbb {N}\),
where \(\mathrm{depth}(g)\) denotes the least depth of a decision tree computing \(g\).
Let \(f\) be a DNF formula on \(n \ge 1\) Boolean variables whose width is at most \(w\), where \(w \ge 1\), and suppose that no term of \(f\) reads any variable more than once: each term is a duplicate-free list of literals, and any two literals in a common term that share a variable are equal. Fix a real parameter \(p\) with \(0 \lt p \le \tfrac {1}{40w}\), so in particular \(p \le 1\). Then for every \(t \in \mathbb {N}\), under a Bernoulli(\(p\)) random restriction \(\rho \) the restricted function \(f\vert _\rho \) (that is, \(f\) evaluated on the free coordinates with the remaining coordinates fixed by \(\rho \)) satisfies
where the decision-tree depth of a Boolean function is the least depth of a decision tree computing it.
Let \(f\) be a DNF formula on \(n \ge 1\) variables in which every term has width at most \(w\), where \(w \ge 1\), and in which each term is reduced: no literal is repeated, and any two literals of a term sharing the same variable are equal (so a variable occurs at most once, and never with both polarities, within a term). For a Boolean function \(g\) on some set of inputs, write \(D(g)\) for its decision-tree depth, the least depth of a decision tree computing \(g\). Fix a natural number \(d\) and an integer \(k\) with \(k \le n\), and draw a restriction \(\rho \) uniformly at random among all restrictions of the \(n\) variables that leave exactly \(k\) variables free (the fixed-size restriction model \(R_k\)). Writing \(f|_\rho \) for the function obtained by fixing each set variable to the value assigned by \(\rho \) and keeping the \(k\) free variables as inputs, one has
Let \(f\) be a DNF formula on \(n \ge 1\) Boolean variables whose width is at most \(w\), with \(w \ge 1\), and suppose that within each term of \(f\) no variable occurs more than once — that is, every term is a repetition-free list of literals in which any two literals sharing a variable are identical. Fix an integer \(k\) with \(0 \le k \le n\), and let \(R_k\) denote the fixed-size restriction model, uniform over all restrictions of the \(n\) variables that leave exactly \(k\) of them free. Then for every \(d\), the probability under \(R_k\) that the restricted function \(f_\rho \) (with the assigned coordinates fixed according to \(\rho \) and the \(k\) free coordinates left as inputs) has decision-tree depth greater than \(d\) satisfies
Let \(f\) be a DNF formula on \(n\) Boolean variables with \(n \ge 1\), and suppose every term of \(f\) has width at most \(w\), so that \(f\) is a disjunction of conjunctions of at most \(w\) literals. Assume that no variable occurs more than once within any single term, that is, each term consists of literals with pairwise distinct variables (and hence has no repeated literal). Fix natural numbers \(k\) and \(d\) with \(5k \le n\), and draw a restriction \(\rho \) uniformly at random from the fixed-size model \(R_k\), which is uniform over all restrictions of the \(n\) variables that leave exactly \(k\) variables free. Then the probability that the restricted function \(x \mapsto f(\rho .\mathrm{extend}\, x)\) has decision-tree depth greater than \(d\) satisfies
where \(\mathrm{depth}(g)\) denotes the least depth of a decision tree computing \(g\).
Let \(f\) be a DNF formula on \(n\) variables whose width is at most \(w\), let \(d\) be a natural number, and let \(\gamma \) be a fixed restriction. Write the Razborov encoding of \(f\) with parameters \(w, d\) on a restriction \(\rho \) as a pair \((\gamma _\rho , \mathrm{aux}_\rho )\), where \(\gamma _\rho \) is the extended restriction and \(\mathrm{aux}_\rho \) is the auxiliary list. Consider those restrictions \(\rho \) that are bad for the Boolean function computed by \(f\) at depth \(d\) — that is, the restriction of that function by \(\rho \) has decision-tree depth strictly greater than \(d\) — and that satisfy \(\gamma _\rho = \gamma \). Then the number of distinct auxiliary lists \(\mathrm{aux}_\rho \) arising from such \(\rho \) is at most \((4w)^d\).
Let \(f\) be a DNF formula on \(n\) Boolean variables in which no two literals of the same term share a variable and no term lists the same literal twice, and suppose every term of \(f\) has at most \(w\) literals. Fix natural numbers \(d \le s\). Then the number of restrictions of the \(n\) variables that leave exactly \(s\) variables free and are bad for \(f\) at depth \(d\) — that is, for which the restricted function \(x \mapsto f(\rho .\mathrm{extend}\, x)\) has decision-tree depth strictly greater than \(d\) — is at most
Let \(f\) be a DNF formula in \(n\) variables, and let \(d\) and \(s\) be natural numbers with \(s \le d\). Then no restriction of the \(n\) variables that leaves exactly \(s\) variables free is bad for the Boolean function computed by \(f\) at depth \(d\); that is, the set of restrictions \(\rho \) that both have exactly \(s\) free variables and whose restricted function \(x \mapsto f(\rho .\mathrm{extend}\, x)\) has decision-tree depth strictly greater than \(d\) is empty.
For a parameter \(p\) and a (decidable) predicate \(\mathrm{event}\) on restrictions, the probability that the predicate holds under a Bernoulli(\(p\)) random restriction is the sum over all restrictions \(\rho \) of \(\mathrm{bernoulliRestrWeight}(p,\rho )\) weighted by the indicator of \(\mathrm{event}(\rho )\).
For a parameter \(p \in \mathbb {R}\) and a restriction \(\rho \) on \(n\) variables, the Bernoulli weight is
so each free variable contributes a factor \(p\) and each fixed variable a factor \((1-p)/2\).
Fix \(n \in \mathbb {N}\) and a real parameter \(p\) with \(0 \le p \le 1\). For a restriction \(\rho \) on \(n\) variables, write \(F(\rho )\) for its set of free variables. Then, summing the Bernoulli weight over all restrictions,
so these weights define a probability distribution on the set of restrictions on \(n\) variables.
The canonical decision tree for \(f|_\rho \), following Razborov’s construction. It is defined through a fuel-driven helper \(\mathrm{canonicalDTree.go}\) initialized with fuel \(\rho .\mathrm{numFree} + \text{(positive)}\), repeatedly selecting the first non-killed term and expanding its free variables.
Fix \(n\) Boolean variables, and let \(f\) be a DNF formula (a list of terms) and \(\rho \) a restriction. Write \(T(f, k, \rho )\) for the fuel-bounded canonical decision tree of \(f\) under \(\rho \) built with fuel budget \(k\), and for a term \(s\), a restriction \(\sigma \), and a continuation \(c\) mapping restrictions to decision trees, write \(S(s, \sigma , c)\) for the term sub-tree of \(s\) under \(\sigma \) with continuation \(c\). Suppose that, under \(\rho \), not every term of \(f\) is killed and no term of \(f\) is fixed, and let \(t\) be the first term of \(f\), in list order, that is not killed by \(\rho \). Then
where the continuation \(c\) sends a restriction \(\rho '\) to the leaf that outputs true when \(t\) is fixed by \(\rho '\), and to \(T(f,\ \mathit{fuel},\ \rho ')\) otherwise.
Let \(f\) be a DNF formula and \(\rho \) a restriction on \(n\) variables such that \(\rho \) does not kill every term of \(f\) and fixes no term of \(f\), and let \(t\) be the first term of \(f\), in list order, that \(\rho \) does not kill. Then the canonical decision tree of \(f\) under \(\rho \) equals the term sub-tree built for \(t\) over \(\rho \) whose continuation, evaluated at each leaf restriction \(\rho '\) reached along a root-to-leaf path, returns the constant-true leaf when \(\rho '\) fixes \(t\) and otherwise returns the fuel-bounded canonical construction for \(f\) under \(\rho '\) run with fuel budget equal to the number of free variables of \(\rho \).
Let \(f\) be a DNF formula on \(n\) Boolean variables and let \(\rho \) be a restriction of those variables. Then the canonical decision tree of \(f\) under \(\rho \) computes exactly the restriction of \(f\) by \(\rho \): for every Boolean assignment \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\), evaluating the tree at \(x\) yields the value of \(f\) on the assignment that takes \(\rho (i)\) wherever \(\rho \) fixes coordinate \(i\) and \(x_i\) on the free coordinates.
Let \(f\) be a DNF formula on \(n\) Boolean variables, let \(\rho \) be a restriction of those variables, and let \(d\) be a natural number. Write \(f|_\rho \) for the restricted Boolean function, which sends an assignment \(x\) to the value of \(f\) on the assignment that agrees with \(\rho \) on its fixed variables and with \(x\) on its free variables. Suppose \(\rho \) is bad for \(f\) at depth \(d\), meaning that the decision-tree depth of \(f|_\rho \) — the least depth of any decision tree computing \(f|_\rho \) — is strictly greater than \(d\). Then the deepest root-to-leaf path of the canonical decision tree of \(f\) under \(\rho \) has length strictly greater than \(d\).
Let \(f\) be a DNF formula on \(n\) Boolean variables and \(\rho \) a restriction, and let \(t\) be the first term of \(f\), in list order, that is not killed by \(\rho \). Suppose the literals of \(t\) are pairwise distinct and no two of them share a variable, and write \(l_{i_0}, l_{i_1}, \ldots \) for the subsequence of literals of \(t\) whose variables are free under \(\rho \), taken in their order of occurrence in \(t\). Then for every index \(k\) smaller than both the number of these free literals and the length of the deepest root-to-leaf path of the canonical decision tree of \(f\) under \(\rho \), the variable queried at the \(k\)-th step of that path equals the variable of the \(k\)-th free literal \(l_{i_k}\).
Let \(f\) be a DNF formula on \(n\) Boolean variables and let \(\rho \) be a restriction on those variables, and write \(f|_\rho \) for the restricted Boolean function \(x \mapsto f\big(\rho .\mathrm{extend}\, x\big)\), which evaluates \(f\) on the assignment that takes the fixed value of \(\rho \) where one is set and reads the remaining coordinates from \(x\). Then the depth of the canonical decision tree for \(f\) under \(\rho \) is at least the decision-tree depth of \(f|_\rho \); that is, the least depth of any decision tree computing \(f|_\rho \) does not exceed the depth of the canonical tree.
Let \(f\) be a DNF formula on \(n\) Boolean variables and let \(\rho \) be a restriction on those variables. If some term of \(f\) is fixed by \(\rho \) — that is, at least one conjunction of literals in \(f\) has every literal fixed to true by \(\rho \) — then the canonical decision tree of \(f\) under \(\rho \) has depth \(0\).
Let \(f\) be a DNF formula on \(n\) Boolean variables and let \(\rho \) be a restriction on those variables. If every term of \(f\) is killed by \(\rho \), then the canonical decision tree for \(f|_\rho \) has depth \(0\).
Let \(f\) be a DNF formula on \(n\) Boolean variables, let \(\rho \) be a restriction of those variables, and let \(m\) be a natural number strictly greater than the number of free variables of \(\rho \). Denote by \(T\) the decision tree obtained from \(f\) and \(\rho \) by the canonical fuel-bounded recursion run with fuel budget \(m\). Then \(T\) computes the restriction of \(f\) by \(\rho \): for every assignment \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\), evaluating \(T\) at \(x\) returns \(f\) applied to the extension of \(x\) along \(\rho \) — the assignment that uses \(\rho \)’s fixed value on each constrained coordinate and \(x\)’s value on each free coordinate.
Let \(f\) be a DNF formula on \(n\) variables, and let \(\rho \) be a restriction on those variables having \(k\) free variables. Write \(G(f, m, \rho )\) for the fuel-bounded recursion that builds a decision tree from the DNF \(f\) and the restriction \(\rho \), in which the natural-number fuel budget \(m\) serves as an upper bound on the recursion depth. Then for any two fuel budgets \(m_1\) and \(m_2\) that both strictly exceed \(k\), the two trees coincide: \(G(f, m_1, \rho ) = G(f, m_2, \rho )\).
Fix \(n\), a DNF formula \(f\) on \(n\) variables, and two restrictions \(\rho _0\) and \(\sigma \) on those variables. Suppose \(t\) is the first term of \(f\) not killed by \(\rho _0\), that \(t\) has no repeated literal, and that any two literals of \(t\) with equal variables are themselves equal. Let \(\mathrm{lits}\) be the list of literals of \(t\) whose variable is free under \(\rho _0\), each tagged with its position in \(t\), and let \(\mathrm{path}\) be a canonical path for \(f|_{\rho _0}\) — an initial segment of the deepest root-to-leaf path of the canonical decision tree of \(f|_{\rho _0}\) — of length at most the depth of the canonical decision tree of \(f|_{\rho _0}\). Assume that for every index \(k\) below the smaller of the lengths of \(\mathrm{lits}\) and \(\mathrm{path}\), the variable of the \(k\)-th literal of \(\mathrm{lits}\) equals the variable queried at the \(k\)-th step of \(\mathrm{path}\). Process \(\mathrm{lits}\) against \(\mathrm{path}\) in lockstep, one path step per literal, forming a restriction \(\rho _0'\) and a shortened path \(\mathrm{path}'\) as follows: a literal with variable \(v\) and negation flag \(b\) aligned with a path step of direction \(d\) contributes the assignment \(v \mapsto d\) to \(\rho _0'\) (and the assignment \(v \mapsto \lnot b\) to \(\sigma \)), and its path step is dropped; \(\rho _0'\) agrees with \(\rho _0\) on every other variable, and \(\mathrm{path}'\) is the unconsumed tail of \(\mathrm{path}\). Then \(\mathrm{path}'\) is a canonical path for \(f|_{\rho _0'}\), and the length of \(\mathrm{path}'\) is at most the depth of the canonical decision tree of \(f|_{\rho _0'}\).
Fix natural numbers \(n\) and \(k\). Among all restrictions on \(n\) variables — maps assigning each index \(i \in \mathrm{Fin}\, n\) either a fixed Boolean value or leaving it free — the number of those having exactly \(k\) free variables equals \(\binom {n}{k}\, 2^{\, n-k}\).
The canonical decision tree of a DNF formula \(f\) under a restriction \(\rho \) is built by a fuel-limited recursion; write \(G_k(f,\rho )\) for the decision tree it returns when run with fuel budget \(k \in \mathbb {N}\) starting from \(\rho \), so that the canonical decision tree of \(f\) under \(\rho \) is \(G_{m(\rho )+1}(f,\rho )\), where \(m(\rho )\) denotes the number of free variables of \(\rho \). Let \(f\) be a DNF formula on \(n\) variables, let \(t\) be a term with \(t \in f\), and let \(\rho _{\mathrm{orig}}\) and \(\rho '\) be restrictions with \(m(\rho _{\mathrm{orig}}) \ge m(\rho ') + 1\). Then the tree that returns the constant-true leaf if \(t\) is fixed by \(\rho '\) and returns \(G_{m(\rho _{\mathrm{orig}})}(f,\rho ')\) otherwise is equal to the canonical decision tree of \(f\) under \(\rho '\).
Fix a number of variables \(n\), and let \(D\) denote the recursive loop underlying Razborov’s decoder, which takes as input a DNF formula \(f\) on \(n\) variables, a width parameter \(w \in \mathbb {N}\), a fuel bound \(\mathrm{fuel} \in \mathbb {N}\), two restrictions \(\sigma \) and \(\rho _0\) on the \(n\) variables, and an auxiliary list of index–value pairs drawn from \(\mathbb {N}\times \mathrm{Bool}\), and returns a pair whose first component is again a restriction on the \(n\) variables. Then for every variable \(v\) with \(\sigma (v) = \mathrm{none}\), the first component of \(D(f, w, \mathrm{fuel}, \sigma , \rho _0, \mathrm{aux})\) also assigns \(\mathrm{none}\) to \(v\).
Let \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be a Boolean function, and let \(T\) be a decision tree on \(n\) Boolean variables that computes \(f\), in the sense that \(T\) evaluates to \(f(x)\) on every input \(x\). Then the depth of \(T\) is at least the decision-tree depth of \(f\); that is, \(\mathrm{depth}(T)\) is no smaller than the least depth of any decision tree computing \(f\).
Let \(f : (\mathrm{Fin}\, n \to \mathbb {F}_2) \to \{ 0,1\} \) be a Boolean function, and let \(d\) be a natural number such that the decision-tree depth of \(f\) is at most \(d\). Then \(f\) is computed by a DNF formula of width at most \(d\) and by a CNF formula of width at most \(d\); that is, there exist a disjunction of terms and a conjunction of clauses, each term or clause containing at most \(d\) literals, that both agree with \(f\) on every input.
Let \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be a Boolean function, and let \(T\) be a decision tree on \(n\) variables of depth at most \(d\) that computes \(f\), meaning \(T\) evaluates to \(f(x)\) on every input \(x\). Then the decision-tree depth of \(f\) is at most \(d\).
Let \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be a Boolean function and let \(\rho \) be a restriction on \(n\) variables, fixing some coordinates to Boolean values and leaving the rest free. Then the decision-tree depth of the restricted function \(f\! \restriction _\rho \), which sends \(x\) to \(f\) evaluated with the free coordinates taken from \(x\) and the remaining coordinates set by \(\rho \), is at most the number of free variables of \(\rho \).
Let \(f \colon (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be a Boolean function on \(n\) variables, and let \(d\) be its decision-tree depth. Then there exists a decision tree \(T\) of depth at most \(d\) that computes \(f\), meaning \(T\) evaluated at \(x\) equals \(f(x)\) for every input \(x\).
Fix a DNF formula \(f\) on \(n\) variables — a list of terms, each a conjunction of literals — together with a width parameter \(w \in \mathbb {N}\), a fuel bound \(\mathrm{fuel} \in \mathbb {N}\), a path \(P\) recording the variable–value pairs already committed (a list of elements of \(\mathrm{Fin}\, n \times \mathrm{Bool}\)), two restrictions \(\rho _0, \sigma : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\), and an accumulator \(A\), a list of index–value pairs in \(\mathbb {N}\times \mathrm{Bool}\). The tail-recursive helper of Razborov’s encoding procedure, run from the initial accumulator \(A\) on these data, produces the same output as the corresponding run started from the empty accumulator, with the entries of \(A\) carried through unchanged; that is, the accumulator contributes to the result only additively and does not affect the encoding computed by the procedure itself.
Consider the recursive encoding routine underlying Razborov’s proof of the switching lemma, which processes a DNF formula \(f\) on \(n\) variables against a width parameter \(w\) and a fuel bound, while carrying a path recorded as a list of variable–value pairs \((i,b) \in \mathrm{Fin}\, n \times \mathrm{Bool}\), two restrictions \(\rho _0, \sigma \) on \(n\) variables, and an accumulator \(\mathrm{acc}\) recorded as a list of index–value pairs in \(\mathbb {N}\times \mathrm{Bool}\); it returns a pair of lists. Then the second component of its output has length at most \(\left\lvert \mathrm{acc}\right\rvert + 2\, \left\lvert \mathrm{path}\right\rvert \), where \(\left\lvert \mathrm{acc}\right\rvert \) and \(\left\lvert \mathrm{path}\right\rvert \) denote the lengths of the accumulator and path lists.
Fix a DNF formula \(f\) on \(n\) variables, a width parameter \(w \in \mathbb {N}\), a fuel bound \(\in \mathbb {N}\), a partial path recorded as a list of variable–value pairs \((i,b) \in \mathrm{Fin}\, n \times \{ 0,1\} \), and two restrictions \(\rho _0, \sigma \) on the \(n\) variables. The auxiliary recursion underlying the Razborov encoding takes, in addition, a starting accumulator: a list of index–bit pairs in \(\mathbb {N}\times \{ 0,1\} \), and returns a pair. Then the first component of its output does not depend on that accumulator: running the recursion from an arbitrary accumulator yields the same first component as running it from the empty accumulator.
Fix a DNF formula \(f\) on \(n\) variables, two restrictions \(\rho _0\) and \(\sigma \), a fuel budget, and a path \(\mathrm{step} :: \mathrm{rest}\) consisting of a first variable–direction pair \(\mathrm{step}\) followed by a list \(\mathrm{rest}\) of further such pairs. Suppose \(t_{\mathrm{clause}}\) is the first term of \(f\) that is not killed by \(\rho _0\), and that, listing the literals of \(t_{\mathrm{clause}}\) each tagged with its position in the term and keeping only those whose variable is free under \(\rho _0\), this filtered list is nonempty, of the form \(\mathit{fl} :: \mathit{fls}\). Then the first component of the encoder run at fuel one greater on the whole path \(\mathrm{step} :: \mathrm{rest}\) coincides with the first component of the encoder run at the smaller fuel started from the state obtained by processing the clause’s free literals \(\mathit{fl} :: \mathit{fls}\) against that same path from \(\rho _0\) and \(\sigma \).
Fix \(n\) and work with formulas on the variables \(\mathrm{Fin}\, n\). Let \(f\) be a DNF formula, \(w \in \mathbb {N}\) a width parameter, and \(\mathrm{fuel} \in \mathbb {N}\) a bound on the recursion depth; let \(\mathrm{path}\) be a list of variable–value pairs (elements of \(\mathrm{Fin}\, n \times \mathrm{Bool}\)), let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables, and let \(\mathrm{acc}\) be a list of index–value pairs holding the encoding accumulated so far. The recursive core of the Razborov encoder sends this data to a pair whose first component is again a restriction; write \(\gamma \) for that restriction. Then for every variable \(v\) that is fixed by \(\rho _0\) — that is, \(\rho _0(v) \neq \mathrm{none}\) — one has \(\gamma (v) = \sigma (v)\).
Fix a DNF formula \(f\) on \(n\) variables, a width \(w\), a fuel bound, a decision path (a list of variable–direction pairs), and a base restriction \(\rho _0\). The recursive worker of the Razborov encoding, run on these inputs together with a running restriction \(\sigma \) and empty auxiliary data, returns among its outputs a restriction; write \(R(\sigma )\) for this first output component. Let \(v\) be a variable that is left free by \(\rho _0\) and by two running restrictions \(\sigma _1\) and \(\sigma _2\), meaning \(\rho _0(v)\), \(\sigma _1(v)\), and \(\sigma _2(v)\) are all undefined. Then the two runs agree at \(v\): \(R(\sigma _1)(v) = R(\sigma _2)(v)\).
Let \(f\) be a DNF formula on \(n\) variables in which no term contains two distinct literals over the same variable, and let \(w\) be a natural-number width parameter. Fix two restrictions \(\rho _0\) and \(\sigma \) on the \(n\) variables such that \(\sigma \) leaves a variable free whenever \(\rho _0\) does (that is, \(\rho _0(v)\) undefined implies \(\sigma (v)\) undefined), and let \(t\) be the first term of \(f\), in list order, that is not killed by \(\rho _0\). Run the recursion at the heart of the Razborov encoding on \(f\) and \(w\) — with any fuel budget, any decision path, initial data \(\rho _0\) and \(\sigma \), and empty accumulated data — and write \(\rho \) for the restriction it returns as its first component. Then for every literal \(l\) of \(t\) whose variable is left free by \(\rho _0\), the output restriction \(\rho \) does not kill \(l\): it does not fix the variable of \(l\) to the value that would make \(l\) false.
Fix a DNF formula \(f\) on \(n\) variables, two natural-number parameters \(w\) and \(t\) (a fuel bound), a list of decisions pairing a variable with a Boolean value, and a restriction \(\rho _0\). Run the Razborov encoding loop on this data, started with empty auxiliary data, using a restriction \(\sigma \). Then the second component of the loop’s output is the same whether one supplies the restriction \(\sigma _1\) or the restriction \(\sigma _2\) in place of \(\sigma \); that is, this component does not depend on the \(\sigma \) argument.
Let \(f\) be a DNF formula on \(n\) variables whose width does not exceed a positive integer \(w\), and fix a recursion-budget (fuel) parameter, a path \(P\) presented as a list of variable–direction pairs in \(\mathrm{Fin}\, n \times \mathrm{Bool}\), and two restrictions \(\rho _0\) and \(\sigma \) on the \(n\) variables. Run the Razborov encoder’s recursive worker on \(f\), \(w\), the fuel, \(P\), \(\rho _0\), and \(\sigma \) starting from an empty auxiliary accumulator, and let \(A\) be its auxiliary output, a list of pairs in \(\mathbb {N}\times \mathrm{Bool}\). Then there exists a list \(ts\) of triples in \(\mathrm{Fin}\, w \times \mathrm{Bool} \times \mathrm{Bool}\) such that \(A\) is exactly the aux list obtained from \(ts\) by the triples-to-aux conversion, the number of triples in \(ts\) is at most the length of \(P\), and, whenever \(ts\) is nonempty, the marker flag (the final Boolean component) of its last triple equals true.
Let \(f\) be a DNF formula on \(n\) variables whose width is at most \(w\), let \(w\) satisfy \(w \gt 0\), let \(d\) be a natural number, and let \(\gamma \) be a fixed restriction on the same \(n\) variables. Consider the collection of all restrictions \(\rho \) that are bad for the Boolean function computed by \(f\) at depth \(d\) — that is, the function \(x \mapsto f(\rho .\mathrm{extend}\, x)\) has decision-tree depth strictly greater than \(d\) — and whose Razborov encoding with parameters \(w,d\) has first component equal to \(\gamma \); write \(S\) for the set of auxiliary lists \(\mathrm{aux} \in (\mathbb {N}\times \mathrm{Bool})^{*}\) arising as the second component of the Razborov encoding of such a \(\rho \). Then there exists a map \(g\) from auxiliary lists to functions \(\mathrm{Fin}\, d \to \mathrm{Fin}\, w \times \mathrm{Bool} \times \mathrm{Bool}\) that is injective on \(S\).
Let \(f\) be a DNF formula on \(n\) variables and let \(\rho \) be a restriction on those variables. Suppose that not every term of \(f\) is killed by \(\rho \), and that no term of \(f\) is fixed by \(\rho \). Then \(\rho \) leaves at least one variable free; that is, there exists an index \(v \in \mathrm{Fin}\, n\) with \(\rho (v)\) unset.
Let \(\rho \) be a restriction on \(n\) variables, let \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\) be a point, and let \(v \in \mathrm{Fin}\, n\) be a free variable of \(\rho \) with \(x(v) = b\). Let \(\rho [v \mapsto b]\) be the restriction obtained from \(\rho \) by fixing the value of variable \(v\) to \(b\) (and leaving every other variable as in \(\rho \)). Then \(\rho [v \mapsto b]\) and \(\rho \) have the same extension along \(x\); that is, \((\rho [v \mapsto b]).\mathrm{extend}\, x\) and \(\rho .\mathrm{extend}\, x\) agree as Boolean assignments.
Let \(f\) be a DNF formula on \(n\) variables whose width is at most \(w\), and suppose no variable occurs twice within any single term of \(f\) (any two literals of a term sharing a variable are equal). Fix natural numbers \(s\) and \(d\) with \(d \le s\), and fix a target restriction \(\gamma \). Consider the restrictions \(\rho \) that are *bad* — having exactly \(s\) free variables while the Boolean function \(x \mapsto f(\rho .\mathrm{extend}\, x)\) computed by \(f\) under \(\rho \) has decision-tree depth strictly greater than \(d\) — and that the Razborov encoding sends to \(\gamma \), i.e. the first component of the encoding of \(f\) and \(\rho \) with parameters \(w, d\) equals \(\gamma \). The number of such restrictions \(\rho \) is at most \((4w)^d\).
Let \(\rho : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\) be a restriction, let \(R\) be a list of literals on the variables \(\mathrm{Fin}\, n\), and let \(v \in \mathrm{Fin}\, n\) and \(b \in \mathrm{Bool}\). Write \(\rho [v \mapsto b]\) for the restriction that agrees with \(\rho \) everywhere except at \(v\), where it takes the fixed value \(b\). If the variable of every literal in \(R\) is distinct from \(v\), then selecting from \(R\) those literals whose variable is free under \(\rho [v \mapsto b]\) yields the same list as selecting those whose variable is free under \(\rho \).
Let \(f\) be a DNF formula on \(n\) variables, viewed as an ordered list of terms, in which no term contains two distinct literals on the same variable. Fix restrictions \(\rho _0\) and \(\sigma \) such that \(\sigma \) leaves free every variable that \(\rho _0\) leaves free, and suppose that \(t_{\mathrm{clause}}\) is the first term of \(f\), in list order, that is not killed by \(\rho _0\). Let \(\rho _{\mathrm{enc}}\) be the restriction obtained as the first component of the output of the Razborov encoding procedure applied to \(f\) — with width parameter \(w\) and the given fuel and path arguments — started from \(\rho _0\) and \(\sigma \), and let \(\rho _{\mathrm{dec}}\) be any restriction that agrees with \(\rho _0\) on every variable that \(\rho _0\) fixes and with \(\rho _{\mathrm{enc}}\) on every variable that \(\rho _0\) leaves free. Then \(t_{\mathrm{clause}}\) is also the first term of \(f\), in list order, that is not killed by \(\rho _{\mathrm{dec}}\).
Let \(f\) be a DNF formula on \(n\) variables — a list of terms, each a conjunction of literals — and let \(\rho \) and \(\sigma \) be restrictions of these \(n\) variables. Suppose \(\sigma \) agrees with \(\rho \) on every variable that \(\rho \) fixes, that is, \(\sigma (v) = \rho (v)\) for every variable \(v\) on which \(\rho \) takes a Boolean value rather than leaving it free. Let \(t\) be the first term of \(f\) (in list order) that is not killed by \(\rho \), and suppose in addition that \(t\) is not killed by \(\sigma \). Then \(t\) is also the first term of \(f\) that is not killed by \(\sigma \).
Let \(f\) be a DNF formula on \(n\) Boolean variables and let \(\rho \) be a restriction of those variables. If at least one term of \(f\) is fixed by \(\rho \), then the restricted Boolean function \(x \mapsto f(\rho \text{-extension of }x)\) has decision-tree depth \(0\).
Fix \(n\), and let \(t\) be a term over \(n\) variables, that is a list of literals, together with a restriction \(\rho _0 : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\), a variable \(v \in \mathrm{Fin}\, n\), and a finite list of entries \((k_1, b_1), \dots , (k_m, b_m)\) with each \(k_i \in \mathbb {N}\) and each \(b_i\) a Boolean. Process these entries from left to right, starting from \(\rho _0\): for an entry \((k_i, b_i)\), delete the first \(k_i\) literals of \(t\); if what remains is nonempty with head literal \(l\), update the current restriction so that the variable of \(l\) is now assigned the fixed value \(b_i\), and otherwise leave the restriction unchanged. Suppose that for every entry \((k_i, b_i)\), whenever deleting the first \(k_i\) literals of \(t\) leaves a nonempty list with head literal \(l\), the variable of \(l\) is not \(v\). Then the restriction obtained after processing all the entries assigns \(v\) the same value as \(\rho _0\) does.
Let \(t\) be a term on \(n\) variables (a list of literals), let \((k_1,b_1),\dots ,(k_m,b_m)\) be a list of pairs in \(\mathbb {N}\times \mathrm{Bool}\), and let \(\rho \) be a restriction on \(n\) variables. Process the entries left to right, starting from \(\rho \): for the entry \((k_i,b_i)\), delete the first \(k_i\) literals of \(t\); if nothing remains, leave the current restriction unchanged, and otherwise reset the variable of the first surviving literal to free. Then for every variable \(v\) left free by \(\rho \), the restriction obtained at the end of this process still leaves \(v\) free.
Fix a term \(t\) over \(n\) variables (a list of literals), a restriction \(\sigma : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\), a variable \(v\), and a finite list of entries, each a pair \((k, b)\) with \(k \in \mathbb {N}\) and \(b\) a Boolean. Process the entries left to right, starting from \(\sigma \): for an entry \((k, b)\), delete the first \(k\) literals of \(t\); if what remains is empty, leave the current restriction unchanged, and otherwise, writing \(l\) for the first remaining literal, reset the restriction to be free (unassigned) at the variable of \(l\). (The Boolean component \(b\) of each entry is never used.) Suppose that for every entry \((k, b)\) in the list, if deleting the first \(k\) literals of \(t\) leaves a nonempty list whose first literal is \(l\), then the variable of \(l\) is not \(v\). Then the restriction obtained after processing all the entries takes the same value at \(v\) as \(\sigma \) does.
Let \(f\) be a DNF formula on \(n\) variables whose width is at most \(w\), and suppose that within every term of \(f\) any two literals naming the same variable are in fact equal. Given a natural number serving as a step budget for the encoder, a finite list of variable–value pairs, and a restriction \(\rho \) on the \(n\) variables, encoding \(\rho \) relative to \(f\) produces a list from which the decoder — run with a step budget one larger than that list’s length — recovers \(\rho \) exactly.
Fix a DNF formula \(f\) on \(n\) variables whose width is at most \(w\) and in which no term repeats a variable (within a single term, any two literals on the same variable are equal). Given a fuel budget, a partial decision path, and restrictions \(\rho _0\) and \(\sigma \), let the encoding procedure, started with empty auxiliary data, return a restriction \(\rho ^\ast \) together with an output list \(c\); let \(D\) denote the decoding procedure run on \(c\) with a fuel budget \(m\) and starting restrictions \(\sigma _{\mathrm{dec}}\) and \(\rho _{0,\mathrm{dec}}\). Suppose that \(\sigma \) is free wherever \(\rho _0\) is free; that on every variable left free by \(\rho _0\) both \(\sigma _{\mathrm{dec}}\) and \(\rho _{0,\mathrm{dec}}\) agree with \(\rho ^\ast \); that on every variable fixed by \(\rho _0\) one has \(\sigma _{\mathrm{dec}} = \sigma \) and \(\rho _{0,\mathrm{dec}} = \rho _0\); and that \(m\) is at least one more than the length of \(c\). Then the restriction returned by \(D\) equals \(\sigma \).
\(\mathrm{IsCanonicalPath}\, f\, \rho \, \mathrm{path}\) holds when \(\mathrm{path}\) is an initial segment of the canonical decision tree’s deepest root-to-leaf path for \(f|_\rho \), i.e. \(\mathrm{path} = (\mathrm{canonicalDTree}\, f\, \rho ).\mathrm{deepPath}.\mathrm{take}\, (\mathrm{path.length})\). This is the invariant carried by the Razborov encoder.
Let \(f\) be a DNF formula on \(n\) Boolean variables and let \(\rho \) be a restriction on these variables. If every term of \(f\) is killed by \(\rho \) — that is, each term contains a literal whose variable \(\rho \) fixes to the value making that literal false — then the restricted function \(x \mapsto f\bigl(\rho .\mathrm{extend}\, x\bigr)\) has decision-tree depth \(0\).
Let \(t\) be a term on \(n\) variables and let \(\rho , \sigma : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\) be two restrictions such that \(\sigma (v) = \rho (v)\) for every variable \(v\) that \(\rho \) fixes (that is, whenever \(\rho (v) \neq \mathrm{none}\)). If \(t\) is killed by \(\rho \), then \(t\) is also killed by \(\sigma \).
Let \(l\) be a list of elements of a type \(\alpha \) and let \(p \colon \alpha \to \{ \mathrm{true}, \mathrm{false}\} \) be a Boolean predicate. If some element \(a\) of \(l\) satisfies \(p(a) = \mathrm{false}\), then the test whether every entry of \(l\) satisfies \(p\) evaluates to \(\mathrm{false}\).
Let \(l\) be a finite list of elements of a type \(\alpha \), and let \(p \colon \alpha \to \{ \mathrm{true}, \mathrm{false}\} \) be a Boolean predicate. If \(p(x) = \mathrm{false}\) for every entry \(x\) of \(l\), then the Boolean existential \(\bigvee _{x \in l} p(x)\) — the value \(\mathrm{true}\) exactly when some entry of \(l\) satisfies \(p\) — is \(\mathrm{false}\).
Let \(l\) be a literal on \(n\) variables and let \(\rho \) be a restriction on those same variables that fixes \(l\), meaning \(\rho \) assigns \(l\)’s variable the Boolean value that makes \(l\) true. Then for every assignment \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\), the literal \(l\) evaluates to true on the extension of \(x\) along \(\rho \) (the assignment using \(\rho \)’s fixed value on each set variable and \(x\) on the free ones).
Fix \(n\) and let \(l\) be a literal on \(n\) variables, with underlying variable \(v \in \mathrm{Fin}\, n\) and polarity flag \(b \in \{ \text{false}, \text{true}\} \) (with \(b = \text{true}\) marking a negated literal). Let \(\rho \) be a restriction on \(n\) variables that kills \(l\), meaning \(\rho \) fixes \(v\) to the value \(b\). Then for every assignment \(x : \mathrm{Fin}\, n \to \{ \text{false}, \text{true}\} \), the literal \(l\) evaluates to false at the extension of \(x\) along \(\rho \).
Let \(w\) be a natural number and let \(b\) be a list of pairs in \(\mathrm{Fin}\, w \times \{ \mathrm{true},\mathrm{false}\} \), and suppose that its marked form — the list of triples obtained by appending to each entry a marker flag that is \(\mathrm{true}\) on the final entry and \(\mathrm{false}\) on every other — is nonempty. Then the marker flag of the last triple of this marked list equals \(\mathrm{true}\).
Fix a width \(w\), and let \(B = ((i_1,b_1),\dots ,(i_k,b_k))\) be a finite list whose entries are pairs, each consisting of an index in \(\{ 0,1,\dots ,w-1\} \) together with a boolean. Form a new list of triples by appending to every entry a marker flag, set to true on the final entry of \(B\) and to false on all earlier entries. Then this list of triples has the same length as \(B\).
Let \(w\) be a natural number and let \(\ell \) be a nonempty list of pairs in \(\mathrm{Fin}\, w \times \{ \mathrm{false},\mathrm{true}\} \). Then the marked list obtained from \(\ell \) by appending a boolean flag to each entry — true on the final entry and false on every other — is again nonempty.
Fix a natural number \(n\), a restriction \(\rho _0\) on \(n\) variables, and a list \(t\) of literals. Enumerate \(t\) as the list of pairs \((\ell , i)\), where \(\ell \) ranges over the literals of \(t\) and \(i\) is the position of \(\ell \) in \(t\), and keep only those pairs whose literal’s variable is a free variable of \(\rho _0\). Then every pair \((\ell , i)\) surviving this filter is itself one of the enumerated pairs of \(t\), and its literal’s variable \(v\) is free under \(\rho _0\), meaning \(\rho _0\) leaves \(v\) unset.
Let \(\rho \) be a restriction on \(n\) variables, let \(v\) be a variable that is free under \(\rho \), and let \(b\) be a Boolean value. Let \(\rho '\) be the restriction obtained from \(\rho \) by fixing \(v\) to the value \(b\) and leaving every other variable unchanged. Then the number of free variables of \(\rho '\) is one less than the number of free variables of \(\rho \); equivalently, if \(f(\rho )\) denotes the number of free variables of \(\rho \), then
Let \(\rho \) be a restriction on \(n\) variables, let \(v \in \mathrm{Fin}\, n\), and let \(b\) be a Boolean value. Write \(\rho [v \mapsto b]\) for the restriction that agrees with \(\rho \) everywhere except at \(v\), where it takes the fixed value \(b\). Then the number of free variables of \(\rho [v \mapsto b]\) is at most the number of free variables of \(\rho \).
Let \(\rho \) be a restriction on \(n\) variables, and let \(v\) be a free variable of \(\rho \). For any Boolean value \(b\), let \(\rho '\) be the restriction that agrees with \(\rho \) everywhere except at \(v\), where it assigns the fixed value \(b\). Then the number of free variables of \(\rho '\) is strictly less than the number of free variables of \(\rho \).
Let \(\rho \) be a restriction on \(n\) variables, let \(v\) be one of these variables, and let \(b\) be a Boolean value. Let \(\rho '\) be the restriction that fixes \(v\) to \(b\) and agrees with \(\rho \) on every other variable. Writing \(f(\sigma )\) for the number of free variables of a restriction \(\sigma \), we have \(f(\rho ') + 1 \ge f(\rho )\); that is, fixing a single variable reduces the number of free variables by at most one.
Parses a low-level aux list of \((\mathbb {N} \times \mathrm{Bool})\) entries into a list of triples \((\mathrm{pos}, \mathrm{dir}, \mathrm{hasMarker})\) with \(\mathrm{pos} : \mathrm{Fin}\, w\), where \(\mathrm{hasMarker}\) records whether the entry is immediately followed by a \((w, \mathrm{false})\) termination marker. Entries with index \(\geq w\) that do not follow a real entry are dropped.
Fix a width \(w \ge 1\) and an index \(\mathrm{idx} \lt w\), together with a direction bit \(\mathrm{dir}\) and an arbitrary list of remaining entries \(\mathrm{rest}\). Parsing the list that begins with the entry \((\mathrm{idx}, \mathrm{dir})\), followed immediately by the termination marker \((w, \mathrm{false})\), and then by \(\mathrm{rest}\), produces the triple \((\mathrm{idx}, \mathrm{dir}, \mathrm{true})\) — with \(\mathrm{idx}\) viewed as an element of \(\{ 0, \dots , w-1\} \) and the final component recording that this entry is marked — prepended to the parse of \(\mathrm{rest}\).
Fix \(w \geq 1\), and consider parsing an aux list whose first two entries are \((\mathrm{idx}, \mathrm{dir})\) and \((\mathrm{idx}', \mathrm{dir}')\), both with in-range indices \(\mathrm{idx} \lt w\) and \(\mathrm{idx}' \lt w\), followed by any further list \(\mathrm{rest}\). Then parsing \((\mathrm{idx}, \mathrm{dir}) :: (\mathrm{idx}', \mathrm{dir}') :: \mathrm{rest}\) yields the triple \((\mathrm{idx}, \mathrm{dir}, \mathrm{false})\) — with no termination marker recorded — prepended to the parse of \((\mathrm{idx}', \mathrm{dir}') :: \mathrm{rest}\).
Fix \(w \ge 1\), and let \((\mathrm{idx}, \mathrm{dir})\) be a single entry whose index satisfies \(\mathrm{idx} \lt w\). Then parsing the one-element list \([(\mathrm{idx}, \mathrm{dir})]\) yields the singleton list \([(\mathrm{idx}, \mathrm{dir}, \mathrm{false})]\), where the index is viewed as an element of \(\{ 0, 1, \dots , w-1\} \) and the marker flag is \(\mathrm{false}\) (there is no following entry to carry a termination marker).
Fix a width \(w \ge 1\), and let \(ts\) be any finite list of triples \((p, d, m)\), where the position \(p\) ranges over \(\{ 0, 1, \dots , w-1\} \) and \(d, m\) are Boolean flags. Converting \(ts\) into an aux list of index–direction entries — appending a termination marker, the pair consisting of the out-of-range index \(w\) together with the flag false, after each triple whose flag \(m\) is true — and then parsing that aux list back into triples recovers exactly the original list \(ts\).
Fix a number \(n\) of variables. Let \(\mathrm{lits}\) be a list of literals, each paired with its position index in a clause; let \(\mathrm{path}\) be a decision-tree path; let \(\rho _0\) and \(\sigma \) be restrictions on the \(n\) variables; and let \(v\) be a variable. Let \(\rho _0'\) denote the updated restriction produced in place of \(\rho _0\) by the clause-processing procedure applied to \(\mathrm{lits}\), \(\mathrm{path}\), \(\rho _0\), and \(\sigma \). If \(\rho _0'\) leaves \(v\) free, then \(\rho _0\) already leaves \(v\) free.
Consumes the free literals of a single clause (each paired with its position index) against the canonical decision-tree path, one path entry per literal. For each free literal it fixes the literal’s variable to its satisfying direction in \(\sigma \), records the path direction in \(\rho _0\) (mirroring the canonical decision-tree branching), and appends the pair (literal position in clause, path direction) to the auxiliary data. It returns the remaining path together with the updated \(\rho _0\), \(\sigma \), and the accumulated clause auxiliary data.
Fix \(n\), and process the free literals of a clause against a decision-tree path: take a list \(\mathrm{lits}\) of literal–index pairs, a decision-tree path, and two restrictions \(\rho _0\) and \(\sigma \) on \(n\) variables, and write \(A\) for the accumulated list of index–direction pairs returned as auxiliary data. Then for every pair \((i, b) \in A\), the index \(i\) equals the index component \(j\) of some literal–index pair \((\ell , j) \in \mathrm{lits}\).
Let \(t\) be a term on \(n\) variables (a list of literals), and let \(\ell \) be a list of literal–index pairs with the property that every pair in \(\ell \) occurs in the indexed enumeration of \(t\); that is, each entry of \(\ell \) is a literal of \(t\) paired with its position in \(t\). Then for any decision-tree path and any restrictions \(\rho _0,\sigma \), every pair \((i,b)\) appearing in the auxiliary output produced when \(\ell \) is processed against the path against \(t\) satisfies \(i \lt \lvert t\rvert \).
Let \(t\) be a term in \(n\) variables, let \(v \in \mathrm{Fin}\, n\) be a variable, and let \(\rho _0, \sigma \) be restrictions on \(n\) variables. Let \(\mathit{lits}\) be a list of literals each tagged with a natural-number position, and suppose that every tagged literal in \(\mathit{lits}\) occurs in the position-indexed enumeration of \(t\) (each literal of \(t\) paired with its index), and that no literal appearing in \(\mathit{lits}\) has variable \(v\). Consider the clause auxiliary data — the list of (position, direction) pairs — produced by processing \(\mathit{lits}\) against a path starting from \(\rho _0\) and \(\sigma \). Then for every pair \((i, b)\) in this auxiliary data, whenever the literals of \(t\) from position \(i\) onward form a nonempty list with leading literal \(\ell \), the variable of \(\ell \) is different from \(v\).
Let \(t\) be a term on \(n\) variables (a list of literals), let \(L\) be a list of literals each paired with a natural-number position, let \(\pi \) be a path (a list of variable–direction pairs), and let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables. Suppose every pair in \(L\) occurs in the indexed enumeration of \(t\), so that each \((l,i)\in L\) records a literal \(l\) sitting at position \(i\) of \(t\). Run the clause-processing procedure that consumes the free literals of \(L\) against \(\pi \) starting from \(\rho _0\) and \(\sigma \), and let \(\rho _0'\) denote its updated \(\rho _0\)-restriction and \(A\) its list of clause auxiliary entries. If \(\rho _0'\) leaves the variable \(v\) free, then for every entry \((i,d)\in A\) and every literal \(l\) such that deleting the first \(i\) literals of \(t\) leaves a nonempty list with head \(l\), the variable of \(l\) is not \(v\).
Let \(t\) be a term on \(n\) variables (a list of literals), let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables, let \(L\) be a list of literals each paired with a position index, and let \(P\) be a decision-tree path. Suppose that every pair \((\ell , i)\) in \(L\) occurs in the indexed enumeration of \(t\) — so that \(\ell \) is precisely the literal at position \(i\) of \(t\) — and that every literal appearing in \(L\) has its variable free under \(\rho _0\) (its value under \(\rho _0\) is \(\mathrm{none}\)). Then for each auxiliary entry \((i, b)\) produced by processing \(L\) against \(P\) starting from \(\rho _0\) and \(\sigma \), and for any literal \(\ell '\) and list \(r\) with \(t\) from position \(i\) onward equal to \(\ell ' :: r\) (i.e. discarding the first \(i\) literals of \(t\) leaves a nonempty list headed by \(\ell '\)), the variable of \(\ell '\) is free under \(\rho _0\).
Fix a list of literal–position pairs \(L\), a path \(P\) (a list of variable–direction pairs), and two restrictions \(\rho _0, \sigma \) on \(n\) variables. Running the clause-processing operation on these inputs yields a remaining path \(P'\) and an auxiliary list \(A\) of position–direction pairs (together with updated restrictions). Then
where \(\left\lvert \cdot \right\rvert \) denotes the number of entries of a list.
Let \(t\) be a term on \(n\) variables (a list of literals), and let \(\mathrm{lits}\) be a list of literal–index pairs, each of which is required to be a literal of \(t\) together with its position — that is, every pair in \(\mathrm{lits}\) occurs in the list obtained by pairing each literal of \(t\) with its index. Fix a path, restrictions \(\rho _0\), \(\sigma \), and \(\rho _0^{\mathrm{dec}}\) on the \(n\) variables, and a variable \(v\) with \(\rho _0(v) = \rho _0^{\mathrm{dec}}(v)\). Processing \(\mathrm{lits}\) against the path starting from \(\rho _0\) and \(\sigma \) yields an updated restriction \(\rho _0'\) together with an auxiliary list of pairs \((i, b)\). Now form a restriction by folding over this auxiliary list, left to right, starting from \(\rho _0^{\mathrm{dec}}\): an entry \((i, b)\) leaves the current restriction unchanged if \(t\) has no literal at position \(i\), and otherwise fixes the variable of the literal at position \(i\) to the Boolean value \(b\). Then the restriction so obtained and \(\rho _0'\) take the same value at \(v\).
Let \(t\) be a term on \(n\) variables, and let \(L\) be a list of literal–index pairs each of which occurs at its recorded position in \(t\) (that is, every pair in \(L\) lies in the enumeration of \(t\) that tags each literal with its index). Fix restrictions \(\rho _0\), \(\sigma \), and \(\rho _0^{\mathrm{dec}}\) on \(n\) variables, a decision-tree path \(P\), and a variable \(v\) that \(\rho _0\) leaves free. Running the clause-processing routine on \(L\), \(P\), \(\rho _0\), \(\sigma \) produces an updated restriction \(\rho _0'\) together with auxiliary data \(A\), a list of (position index, direction) pairs, and assume that \(\rho _0'\) assigns a value at \(v\). Consider the restriction obtained by folding \(A\) over the starting restriction \(\rho _0^{\mathrm{dec}}\), where each pair \((i, b)\) leaves the current restriction unchanged if \(t\) has no literal at position \(i\), and otherwise sets the variable of the \(i\)-th literal of \(t\) to the value \(b\). Then this folded restriction agrees with \(\rho _0'\) at \(v\); that is, its value at \(v\) equals \(\rho _0'(v)\).
Let \(t\) be a term on \(n\) variables — a list of literals, each carrying a variable and a negation flag — and refer to the position of a literal by its index in this list. Let \(\mathrm{lits}\) be a list of (literal, index) pairs, every one of which is the literal occurring at that index in \(t\); let \(\mathrm{path}\) be a list of (variable, Boolean) pairs; and let \(\rho _0\), \(\sigma \), and \(\sigma _{\mathrm{dec}}\) be restrictions on \(n\) variables, with \(v\) a distinguished variable. Feeding \(\mathrm{lits}\), \(\mathrm{path}\), \(\rho _0\), and \(\sigma \) to the clause-processing procedure returns an updated restriction \(\rho _0'\) together with an auxiliary list \(A\) of (index, direction) pairs. Suppose \(\rho _0(v)\) is free while \(\rho _0'(v)\) is fixed. Then processing \(A\) from left to right, starting from \(\sigma _{\mathrm{dec}}\) and, at each entry whose index component is \(i\), resetting to free the variable of the \(i\)-th literal of \(t\) whenever \(t\) has a literal at position \(i\) (and otherwise leaving the restriction unchanged), yields a restriction that is free at \(v\).
Let \(t\) be a term on \(n\) variables (a list of literals) and let \(\rho \) be a restriction on those variables. Suppose that \(t\) contains no repeated literal, and that any two literals of \(t\) sharing the same underlying variable are in fact equal. Form the list of literals of \(t\) whose variable is free under \(\rho \), each kept together with its position in \(t\). Then this list has pairwise distinct variables: any two of its entries carry different variables.
Fix a list of literals with position indices and a decision-tree path, and let \(\rho _0\) and \(\sigma \) be two restrictions on \(n\) variables that agree on which variables are free: for every variable \(v\), \(\rho _0(v)\) is free if and only if \(\sigma (v)\) is free. Let \(\rho _0'\) and \(\sigma '\) be the updated restrictions produced by processing the clause’s literals against the path starting from \(\rho _0\) and \(\sigma \). Then \(\rho _0'\) and \(\sigma '\) again agree on their free variables: for every \(v\), \(\rho _0'(v)\) is free if and only if \(\sigma '(v)\) is free.
Let \(\ell \) be a list of free literals of a clause, each paired with its position index, and let \(p\) be a decision-tree path, that is, a list of variable–direction pairs; let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables. Then the remaining path returned by processing \(\ell \) against \(p\) (starting from \(\rho _0\) and \(\sigma \)) is exactly the suffix of \(p\) obtained by deleting its first \(\min (|\ell |,\, |p|)\) entries.
Let \(n\) be a natural number, let \(\mathit{lits}\) be a list of literals on \(n\) variables each paired with its position index, let \(\mathit{path}\) be a list of variable–direction pairs, and let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables. Running the clause-processing procedure on \(\mathit{lits}\), \(\mathit{path}\), \(\rho _0\), and \(\sigma \) yields, among its outputs, a remaining path and a list of accumulated auxiliary pairs. The length of the auxiliary list plus the length of the remaining path is at most the length of \(\mathit{path}\).
Fix a number \(n\) of variables, and let \(\rho _0\) and \(\sigma \) be restrictions on them. Suppose the free literals of a clause are given as a list, each literal paired with its position index, and let a decision-tree path be a list of variable–direction pairs whose length is at least the number of literals. Running the clause-processing operation on this list, this path, \(\rho _0\), and \(\sigma \) returns an updated restriction \(\rho _0'\) in place of \(\rho _0\). If a variable \(v\) is free under \(\rho _0\) and remains free under \(\rho _0'\), then no literal on the list has \(v\) as its underlying variable.
Let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables, and write \(f(\rho )\) for the number of free variables of a restriction \(\rho \). Present a single clause as a list of \(k\) literals, each paired with its position index, and a decision-tree path as a list of \(m\) variable–direction pairs. Suppose that every literal’s variable is free in \(\rho _0\) and that the listed literals have pairwise distinct variables. Let \(\rho _0'\) be the updated restriction returned in place of \(\rho _0\) when these literals are processed against the path. Then
so processing fixes exactly \(\min (k, m)\) of the previously free variables of \(\rho _0\).
Fix a list of literals on \(n\) variables, each paired with its position index within its clause; a decision path presented as a list of variable–direction pairs; and two restrictions \(\rho _0, \sigma \) on \(n\) variables. For a restriction \(\rho \), write \(F(\rho )\) for the number of variables left free by \(\rho \). Suppose that \(\rho _0\) and \(\sigma \) are free at exactly the same variables (for every variable \(v\), \(\rho _0(v)\) is unset if and only if \(\sigma (v)\) is unset), that each listed literal’s variable is free in \(\rho _0\), and that the listed literals have pairwise distinct variables. Let \(\sigma '\) be the restriction returned as the \(\sigma \)-component of processing these literals against the path. Then
where \(\ell \) is the number of literals and \(m\) is the length of the path.
Let \(n\) be a natural number, let \(\mathit{lits}\) be a list of literals on \(n\) variables each tagged with its position index, let \(\mathit{path}\) be a list of decision-tree branch entries (each a variable together with a Boolean direction), and let \(\rho _0, \sigma \) be restrictions on \(n\) variables. Then the remaining path returned by processing \(\mathit{lits}\) against \(\mathit{path}\) starting from \(\rho _0\) and \(\sigma \) — that is, the first component of the result of this processing — has length at most the length of \(\mathit{path}\).
Run the clause-processing operation on a list of literals — each tagged with its position in the clause — and a decision-tree path (a list of variable–direction pairs), starting from restrictions \(\rho _0\) and \(\sigma \). Writing \(\ell \) for the number of literals and \(m\) for the length of the path, the residual path returned by the operation has length exactly \(m - \min (\ell , m)\); equivalently, this residual length together with \(\min (\ell , m)\) sums to \(m\).
Fix \(n\), a list \(\mathit{lits}\) whose entries are literals paired with clause positions, a path \(\mathit{path}\) presented as a list of variable–direction pairs, and two restrictions \(\rho _0,\sigma \) on \(n\) variables. Suppose the pair \((l,\mathit{idx})\) occurs in \(\mathit{lits}\), that every entry of \(\mathit{lits}\) whose variable equals \(\mathrm{var}(l)\) is equal to \(l\) itself, and that \(\rho _0\) leaves \(\mathrm{var}(l)\) free. Processing \(\mathit{lits}\) against \(\mathit{path}\) from the starting restrictions \(\rho _0\) and \(\sigma \) yields an updated restriction \(\rho _0'\), an updated restriction \(\sigma '\), a remaining path, and accumulated auxiliary data. If \(\rho _0'\) still leaves \(\mathrm{var}(l)\) free, then the remaining path yielded is empty.
Let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables, let \(L\) be a list of literals each paired with its position index in a clause, and let \(P\) be a decision-tree path (a list of variable–direction pairs). Suppose the length of \(P\) is at most the number of free variables of \(\rho _0\). Write \((P', \rho _0', \sigma ', A)\) for the result of processing \(L\) against \(P\) starting from \(\rho _0\) and \(\sigma \), so that \(P'\) is the remaining path and \(\rho _0'\) the updated first restriction. Then the length of \(P'\) is at most the number of free variables of \(\rho _0'\).
Fix \(n\), and let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables, let \(\ell \) be a list of literals each paired with its position index in a clause, and let a decision-tree path be given. Suppose the variable \(v\) is fixed under \(\rho _0\), i.e. \(\rho _0(v) \neq \mathrm{none}\). Then, after processing the clause’s free literals against the path — which returns updated versions of \(\rho _0\), \(\sigma \), together with the remaining path and the accumulated auxiliary data — the variable \(v\) remains fixed in the returned restriction \(\rho _0\): its value there is again not \(\mathrm{none}\). In other words, this process never frees a variable that was already fixed in \(\rho _0\).
Let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables, let a list of literals be given (each paired with its position index inside its clause), and let a decision-tree path be given as a list of variable–direction pairs, with at least as many path entries as literals in the list. Process the literals against the path starting from \(\rho _0\) and \(\sigma \), and let \(\rho _0'\) be the updated first restriction that this returns. If some literal in the list has underlying variable \(v\), then \(v\) is not free in \(\rho _0'\); that is, \(\rho _0'(v)\) is a fixed Boolean value rather than the free marker.
Fix \(n\) and let \(\mathrm{lits}\) be a list of literals on \(n\) variables, each paired with a position index, and let \(\mathrm{path}\) be a list of variable–direction pairs. Given restrictions \(\rho _0,\sigma \) on \(n\) variables and a variable \(v\) such that \(v\) is not the variable of any literal occurring in \(\mathrm{lits}\), the updated restriction \(\rho _0'\) produced by processing the clause’s free literals of \(\mathrm{lits}\) against \(\mathrm{path}\) starting from \(\rho _0\) and \(\sigma \) agrees with \(\rho _0\) at \(v\); that is, \(\rho _0'(v) = \rho _0(v)\).
Consider the routine that processes one clause’s free literals — a list of literals, each tagged with its position index in the clause — against a canonical decision-tree path, starting from a restriction \(\rho _0\) and a restriction \(\sigma \), and returning (among other data) an updated restriction \(\sigma '\). Fix the list of literals, the path, the restriction \(\rho _0\), and a variable \(v\). If two starting restrictions \(\sigma _1\) and \(\sigma _2\) agree at \(v\), that is \(\sigma _1(v) = \sigma _2(v)\), then the updated restrictions produced from \(\sigma _1\) and from \(\sigma _2\) also agree at \(v\); equivalently, the value at \(v\) of the updated \(\sigma \)-restriction depends only on the starting value \(\sigma (v)\).
Fix a list of literals, each paired with its position index within a clause, a decision-tree path, a restriction \(\rho _0\), and two arbitrary restrictions \(\sigma _1\) and \(\sigma _2\) on \(n\) variables. Running the clause-processing procedure on this data with \(\sigma _1\) and with \(\sigma _2\) produces the same three outputs in each case: the same remaining path, the same updated restriction \(\rho _0\), and the same accumulated auxiliary list of (position index, path direction) pairs. In other words, of the procedure’s outputs only the returned \(\sigma \) can depend on the \(\sigma \) supplied as input; the path, the \(\rho _0\) component, and the auxiliary data are all unaffected by that choice.
Let \(\ell \) be a literal with variable \(v\) and negation flag \(b\), let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables, and let \(\mathit{lits}\) be a list of literals (each tagged with its position index in the clause) run against a decision-tree path. Suppose that every literal in \(\mathit{lits}\) whose variable is \(v\) is equal to \(\ell \), and that \(\sigma \) does not already assign \(v\) the fixed value \(b\), i.e. \(\sigma (v) \neq \mathrm{some}\, b\). Then the restriction \(\sigma '\) obtained as the \(\sigma \)-component of processing \(\mathit{lits}\) against the path starting from \(\rho _0\) and \(\sigma \) still does not assign \(v\) the value \(b\), i.e. \(\sigma '(v) \neq \mathrm{some}\, b\).
Fix \(n\), and let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables. Run the clause-processing procedure on a list of literals (each paired with its position index) together with a decision-tree path, starting from \(\rho _0\) and \(\sigma \), and write \(\rho _0'\) and \(\sigma '\) for the two restrictions it returns. If a variable \(v\) is free under the input, \(\rho _0(v) = \mathrm{none}\), and remains free under the output, \(\rho _0'(v) = \mathrm{none}\), then the procedure has not altered \(\sigma \) at \(v\) either: \(\sigma '(v) = \sigma (v)\).
Fix \(n\) and let \(\mathrm{lits}\) be a list of literals over \(\mathrm{Fin}\, n\), each paired with a position index, let \(\mathrm{path}\) be a list of variable–direction pairs, and let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables. Suppose \(v \in \mathrm{Fin}\, n\) is not the variable of any literal occurring in \(\mathrm{lits}\). Then the restriction \(\sigma '\) produced by processing \(\mathrm{lits}\) against \(\mathrm{path}\) starting from \(\rho _0\) and \(\sigma \) agrees with \(\sigma \) at \(v\), that is \(\sigma '(v) = \sigma (v)\).
Let \(t\) be a term on \(n\) variables whose literals have pairwise distinct variables, let \(\rho _0\) and \(\sigma \) be restrictions on \(n\) variables, and let \(c\) be a continuation sending each restriction to a decision tree. Let \(S\) be the sub-tree built for \(t\) under \(\rho _0\) with continuation \(c\), and let \(\pi \) denote its deepest root-to-leaf path. Write \(F\) for the sublist of \(t\) consisting of those literals whose variable is free under \(\rho _0\), and set \(m = |F|\). Suppose \(L\) is a list of \(m\) literal–index pairs whose literal components are exactly \(F\), in order, and that \(P\) is a list of variable–direction pairs coinciding with the length-\(|P|\) prefix of \(\pi \), with \(m \le |P| \le |\pi |\). Then dropping the first \(m\) entries of \(\pi \) leaves exactly the deepest path of \(c\) evaluated at the restriction obtained from \(\rho _0\) and \(\sigma \) by processing the literals \(L\) against the path \(P\) (the updated \(\rho _0\) component that this process returns).
Run the clause-processing procedure on a nonempty list of free literals — each paired with its position index in the clause — and a nonempty canonical path, starting from restrictions \(\rho _0\) and \(\sigma \), and let its output consist of a remaining path and a list of accumulated auxiliary entries (among the updated restrictions). Writing \(a\) for the number of auxiliary entries produced, \(r\) for the length of the remaining path returned, and \(m\) for the length of the input path, we have
Work over \(n\) Boolean variables, where a *literal* is a variable together with a negation flag, a *term* is a list of literals (a conjunction), and a *restriction* \(\rho : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\) fixes some variables and leaves the rest free. Let \(t\) be a term whose length is at most a bound \(w\), let \(\mathit{lits}\) be a list of literals each paired with a position index, every such pair occurring among the indexed entries of \(t\), and fix a canonical decision-tree path together with restrictions \(\rho _0^{\mathrm{enc}}, \sigma ^{\mathrm{enc}}, \sigma ^{\mathrm{dec}}, \rho _0^{\mathrm{dec}}\) and a remainder list. Then, running the decoder’s entry-processing routine on the auxiliary data emitted when this clause’s free literals are consumed against the path — followed by the termination marker \((w, \mathrm{false})\) and the remainder — yields the same \(\sigma \)- and \(\rho _0\)-restrictions as folding the corresponding per-entry updates directly over that auxiliary data, with the remainder left unconsumed.
Fix \(n\) variables and consider the entry-processing step of the Razborov decoding procedure, which takes a term \(t\) (a conjunction of literals) over these variables, a natural number \(w\), two restrictions \(\sigma , \rho _0 \colon \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\), and a list \(\mathit{aux}\) of pairs \((\mathbb {N} \times \mathrm{Bool})\), and returns a restriction as the first component of its output. If \(\sigma \) leaves a variable \(v\) free (that is, \(\sigma (v) = \mathtt{none}\)), then the restriction produced by this step also leaves \(v\) free.
The inverse of the Razborov encoding: given the DNF \(f\), parameter \(w\), and a pair \((\gamma , \mathrm{aux})\) produced by the encoder, it recovers the original restriction \(\rho \). It repeatedly finds the first clause not killed by the current restriction, processes that clause’s auxiliary block (unfixing each recorded variable in \(\sigma \) and recording its path direction in \(\rho _0\) until a termination marker is reached), and returns the resulting restriction.
Let \(f\) be a DNF formula on \(n\) variables whose width is at most \(w\), and suppose that no term of \(f\) contains two distinct literals on the same variable (within each term, any two literals sharing a variable coincide). Let \(\rho \) be a restriction that is bad for \(f\) at depth \(d\), meaning the restriction of the Boolean function computed by \(f\) along \(\rho \) has decision-tree depth strictly greater than \(d\). Writing \((\gamma , \mathrm{aux})\) for the pair produced by the Razborov encoding of \(f\), \(w\), \(d\), \(\rho \), the Razborov decoding of \((\gamma , \mathrm{aux})\) with parameters \(f\) and \(w\) returns the original restriction \(\rho \).
The Razborov encoding of a DNF \(f\) and a restriction \(\rho \) with parameters \(w, d\). It takes the first \(d\) steps of the deepest path of the canonical decision tree for \(f|_\rho \), then repeatedly selects the first clause of \(f\) not killed by the current restriction, processes all of that clause’s free literals against the path, and emits a termination marker \((w, \mathrm{false})\). The result is a pair \((\gamma , \mathrm{aux})\), where \(\gamma \) extends \(\rho \) by fixing \(d\) variables to their satisfying directions and \(\mathrm{aux}\) consists of per-clause blocks separated by the termination markers.
Let \(f\) be a DNF formula on \(n\) Boolean variables, let \(w, d \in \mathbb {N}\), and let \(\rho \) be a restriction that is bad for \(f\) at depth \(d\); that is, the restriction of the Boolean function computed by \(f\) along \(\rho \) has decision-tree depth strictly greater than \(d\). Then the auxiliary component of the Razborov encoding of \(f\) and \(\rho \) with parameters \(w\) and \(d\) — the list of \((\mathbb {N}\times \mathrm{Bool})\) pairs returned as the second entry of the encoding — has length at most \(2d\).
Let \(f\) be a DNF formula on \(n\) Boolean variables in which no term repeats a literal, and within each term any two literals on the same variable coincide (so each variable occurs at most once per term). Let \(\rho \) be a restriction with exactly \(s\) free variables, and let \(w\) and \(d\) be natural numbers with \(d \le s\). Suppose \(\rho \) is bad for \(f\) at depth \(d\), meaning the restricted Boolean function \(f|_\rho \) has decision-tree depth strictly greater than \(d\). Then the restriction forming the first component of the Razborov encoding of \(f\) and \(\rho \) with parameters \(w\) and \(d\) has exactly \(s - d\) free variables.
Let \(f\) be a DNF formula on \(n\) Boolean variables in which, within each term, any two literals on the same variable are equal and no literal occurs more than once, and fix an arbitrary width \(w \in \mathbb {N}\). Let \(\rho _0\) and \(\sigma \) be restrictions on these \(n\) variables that leave exactly the same variables free (for every variable \(v\), \(\rho _0(v)\) is unset iff \(\sigma (v)\) is), and let \(\mathrm{path}\) be a prefix of the deepest root-to-leaf path of the canonical decision tree for \(f|_{\rho _0}\), of length at most the depth of that tree. If the fuel parameter strictly exceeds the length of \(\mathrm{path}\), then the restriction returned by the Razborov encoder run on \(f\) with width \(w\), this fuel, \(\mathrm{path}\), \(\rho _0\), and \(\sigma \) has exactly \(\lvert \mathrm{path}\rvert \) fewer free variables than \(\sigma \); that is, its number of free variables plus \(\lvert \mathrm{path}\rvert \) equals the number of free variables of \(\sigma \).
Fix a DNF formula \(f\) on \(n\) variables and parameters \(w, d \in \mathbb {N}\), and suppose that \(f\) has width at most \(w\) and that no term of \(f\) repeats a variable — within any single term, two literals sharing the same variable are identical. Under these hypotheses the Razborov encoding with parameters \(w\) and \(d\) is injective on the restrictions that are bad for \(f\) at depth \(d\): if \(\rho _1\) and \(\rho _2\) are restrictions, each bad for the Boolean function computed by \(f\) at depth \(d\) (that is, the corresponding restricted function has decision-tree depth exceeding \(d\)), and their Razborov encodings agree, then \(\rho _1 = \rho _2\).
Let \(f\) be a DNF formula on \(n\) variables, let \(\rho \) be a restriction of those variables, and let \(d\) be a natural number. Suppose \(\rho \) is bad for \(f\) at depth \(d\): the restriction \(x \mapsto f(\rho .\mathrm{extend}\, x)\) of the Boolean function computed by \(f\) has decision-tree depth strictly greater than \(d\). Then the deepest root-to-leaf path of the canonical decision tree of \(f\) under \(\rho \), truncated to its first \(d\) steps, is a sequence of exactly \(d\) pairs (queried variable, branch direction) — equivalently, that path has length at least \(d\).
Given \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) and a restriction \(\rho \), the restricted function maps \(x\) to \(f(\rho .\mathrm{extend}\, x)\), i.e. \(f\) evaluated on the free coordinates with the fixed coordinates set by \(\rho \).
Given a restriction \(\rho \) and a point \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\), the extension \(\rho .\mathrm{extend}\, x\) is the Boolean assignment that uses the fixed value \(\rho (i)\) where it is set and falls back to \(x(i)\) on the free variables.
Fix \(n\), and let \(f\) be a DNF formula on \(n\) variables and \(w\) a natural number. Let \(\rho _0\), \(\sigma \), \(\sigma _{\mathrm{dec}}\), and \(\rho _0^{\mathrm{dec}}\) be restrictions on \(n\) variables (each variable either assigned a Boolean value or left free), and let \(d\) be a fuel bound. Suppose that \(\sigma \) leaves free every variable that \(\rho _0\) leaves free, and that \(\sigma _{\mathrm{dec}}\) agrees with \(\sigma \) on every variable — both those that \(\rho _0\) leaves free and those that \(\rho _0\) fixes. Then running the Razborov decoding loop on \(f\), with width \(w\), fuel \(d\), working restrictions \(\sigma _{\mathrm{dec}}\) and \(\rho _0^{\mathrm{dec}}\), and the empty list of emitted advice, yields an output whose first component is exactly \(\sigma \).
Let \(t\) be a term on \(n\) variables, let \(\mathit{lits}\) be a list of literals each tagged with a position index, and let \(\mathit{path}\) be a list of variable–direction pairs. Let \(\rho _0\), \(\sigma \), and \(\sigma _{\mathrm{dec}}\) be restrictions on \(n\) variables, and suppose that every variable occurring in a literal of \(\mathit{lits}\) is free under \(\rho _0\); that every tagged literal in \(\mathit{lits}\) occurs at its recorded position in \(t\); that \(\sigma \) is free at every variable that is free under \(\rho _0\); and that \(\sigma _{\mathrm{dec}}\) agrees with \(\sigma \) at every variable that \(\rho _0\) fixes. Write \((\mathit{path}', \rho _0', \sigma ', a)\) for the result of processing \(\mathit{lits}\) along \(\mathit{path}\) starting from \(\rho _0\) and \(\sigma \), where \(a\) is the accumulated list of (position, direction) pairs. Then for every variable \(v\) that \(\rho _0'\) fixes, the restriction obtained by folding over \(a\) — starting from \(\sigma _{\mathrm{dec}}\) and, for each recorded pair with position \(i\), freeing the variable of the literal at position \(i\) in \(t\) (and leaving the restriction unchanged when \(i\) lies beyond the end of \(t\)) — agrees with \(\sigma \) at \(v\).
Let \(t\) be a term on \(n\) variables (a list of literals), let \(\mathit{lits}\) be a list of literals each paired with a position index, let a decision-tree path be given, and let \(\rho _0\), \(\sigma \), and \(\rho _{0,\mathrm{dec}}\) be restrictions on the \(n\) variables. Assume that every literal in \(\mathit{lits}\) has its variable left free by \(\rho _0\), that every pair in \(\mathit{lits}\) occurs in the indexed enumeration of \(t\) (each literal of \(t\) paired with its position), and that \(\rho _{0,\mathrm{dec}}\) agrees with \(\rho _0\) at every variable that \(\rho _0\) fixes. Write \(\rho _0^{\ast }\) for the restriction and \(A\) for the auxiliary list of index–direction pairs produced by processing \(\mathit{lits}\) against the path starting from \(\rho _0\) and \(\sigma \). Reconstruct a restriction from \(\rho _{0,\mathrm{dec}}\) by folding over \(A\): for each pair \((i,b)\), if deleting the first \(i\) literals of \(t\) leaves a nonempty list, set the variable of that list’s head literal to \(b\), and otherwise leave the restriction unchanged. Then at every variable \(v\) at which \(\rho _0^{\ast }\) is defined, this reconstructed restriction and \(\rho _0^{\ast }\) take the same value.
Let \(f\) be a DNF formula on \(n\) variables and \(\rho \) a restriction. Suppose that at least one term of \(f\) is not killed by \(\rho \), and that no term of \(f\) is fixed by \(\rho \). Then the branch-variable selection procedure applied to \(f\) and \(\rho \) succeeds, returning a variable \(v\) that is free in \(\rho \).
Let \(f\) be a DNF formula on \(n \ge 1\) variables in which no variable occurs more than once within any single term, and let \(w, s \in \mathbb {N}\) satisfy \(5s \le n\) and \(\mathrm{width}(f) \le w\) (that is, every term of \(f\) has at most \(w\) literals). For a restriction \(\rho \) that fixes some variables and leaves the others free, write \(f\! \restriction _\rho \) for the Boolean function obtained by evaluating \(f\) with the fixed coordinates set by \(\rho \) and the free coordinates supplied as input. Then the number of restrictions \(\rho \) leaving exactly \(s\) of the \(n\) variables free for which \(f\! \restriction _\rho \) agrees with no CNF formula of width at most \(w\) satisfies
Let \(n \ge 1\), and let \(f\) be a DNF formula on \(n\) Boolean variables whose width is at most \(w\), so that every term of \(f\) has at most \(w\) literals. Suppose that \(5s \le n\), that no term of \(f\) contains two distinct literals on the same variable, and that no term of \(f\) repeats a literal. Call a restriction \(\rho \) on the \(n\) variables an \(s\)-restriction if it leaves exactly \(s\) variables free, and call it bad (at depth \(d\)) if the restricted function \(f|_\rho \) has decision-tree depth strictly greater than \(d\). Then the number \(N\) of \(s\)-restrictions that are bad satisfies
Given a term \(t\) (a conjunction of literals) and a restriction \(\rho \), this is the list of literals of \(t\) whose underlying variable is still free under \(\rho \), i.e. the sublist of \(t\) obtained by keeping those literals \(l\) with \(l.\mathrm{var} \in \rho .\mathrm{freeVars}\).
Let \(f\) be a DNF formula in \(n\) variables, that is, a disjunction \(t_1 \vee \cdots \vee t_m\) of terms, where each term is a conjunction of literals, and let its width be the largest number of literals occurring in any of its terms (and \(0\) if \(f\) has no terms). Then for every term \(t\) appearing in \(f\), the number of literals in \(t\) is at most the width of \(f\).
Builds a complete sub-tree for a term, queried as a list of literals: it branches on each free variable of the term in order, and at every leaf invokes a continuation \(\mathrm{cont} : \mathrm{Restriction}\, n \to \mathrm{DecisionTree}\, n\) with the restriction updated along that root-to-leaf path. Non-free literals are skipped.
Let \(l\) be a literal on \(n\) Boolean variables, write \(i\) for its underlying variable, let \(\mathit{rest}\) be a further list of literals, let \(\rho \) be a restriction on these \(n\) variables, and let the continuation assign a decision tree to each restriction. If \(i\) is free under \(\rho \), then the sub-tree built for the list \(l :: \mathit{rest}\) under \(\rho \) (with this continuation) is the branch that queries \(i\), whose false-child is the sub-tree built for \(\mathit{rest}\) under the restriction obtained from \(\rho \) by fixing \(i\) to false, and whose true-child is the sub-tree built for \(\mathit{rest}\) under the restriction obtained from \(\rho \) by fixing \(i\) to true, both with the same continuation.
Let \(\rho \) be a restriction on \(n\) variables, let \(l\) be a literal whose underlying variable \(v\) is not free under \(\rho \) (that is, \(\rho \) already assigns \(v\) a fixed Boolean value), let \(\mathit{rest}\) be a list of literals, and let \(\mathrm{cont} : \mathrm{Restriction}\, n \to \mathrm{DecisionTree}\, n\) be a continuation. Then the sub-tree built for the term \(l :: \mathit{rest}\) under \(\rho \) with continuation \(\mathrm{cont}\) coincides with the one built for \(\mathit{rest}\) under \(\rho \) with the same continuation; the fixed head literal \(l\) contributes no branch.
Let \(\ell \) be a list of literals on \(n\) variables and let \(\rho \) be a restriction. For a continuation \(c : \mathrm{Restriction}\, n \to \mathrm{DecisionTree}\, n\), write \(T(\ell , \rho , c)\) for the term sub-tree built from \(\ell \) starting at \(\rho \) with continuation \(c\). If two continuations \(c_1\) and \(c_2\) agree on every restriction \(\rho '\) whose number of free variables is at most the number of free variables of \(\rho \) — that is, \(c_1(\rho ') = c_2(\rho ')\) whenever \(\rho '\) has no more free variables than \(\rho \) — then \(T(\ell , \rho , c_1) = T(\ell , \rho , c_2)\).
Let \(L\) be a list of literals on \(n\) variables and \(\rho \) a restriction on those variables, and suppose at least one literal in \(L\) has its variable free under \(\rho \). Let \(c_1\) and \(c_2\) be two continuations — maps sending each restriction to a decision tree — that agree on every restriction having strictly fewer free variables than \(\rho \). Then the two term sub-trees built from \(L\) and \(\rho \), one using the continuation \(c_1\) and the other using \(c_2\), are identical.
Let \(\rho \) be a restriction on \(n\) variables, let \(\mathit{lits}\) be a list of literals whose underlying variables are pairwise distinct, and let \(\mathrm{cont}\) be a continuation assigning a decision tree to every restriction. Then there is a restriction \(\rho '\) that agrees with \(\rho \) at every variable not occurring among the literals of \(\mathit{lits}\), and such that the deepest root-to-leaf path of the term sub-tree built from \(\mathit{lits}\), \(\rho \), and \(\mathrm{cont}\) has length equal to the number of literals in \(\mathit{lits}\) whose variable is free under \(\rho \), plus the length of the deepest root-to-leaf path of \(\mathrm{cont}(\rho ')\).
Fix a restriction \(\rho \) on \(n\) variables and a continuation \(c\) that assigns a decision tree to each restriction. Let \(l\) be a literal whose variable \(v\) is free in \(\rho \), and let \(\mathit{rest}\) be a further list of literals. Then there is a Boolean value \(b\) such that the deepest root-to-leaf path of the term sub-tree built from the literal list \(l :: \mathit{rest}\), the restriction \(\rho \), and the continuation \(c\) begins with the pair \((v, b)\) and continues with the deepest root-to-leaf path of the term sub-tree built from \(\mathit{rest}\), the restriction \(\rho \) with \(v\) now fixed to the value \(b\), and the same continuation \(c\).
Let \(\mathit{lits}\) be a list of literals on \(n\) variables in which no two literals share the same variable, let \(\rho \) be a restriction on \(n\) variables, and let \(c\) be a continuation, that is, a map sending each restriction to a decision tree. Then there exist a list \(P\) of (variable, branch-direction) pairs and a restriction \(\rho '\) such that the deepest path of the sub-tree for the term \(\mathit{lits}\) built from \(\rho \) and \(c\) equals \(P\) followed by the deepest path of \(c(\rho ')\); the length of \(P\) equals the number of literals of \(\mathit{lits}\) whose variable is free under \(\rho \); and \(\rho '\) agrees with \(\rho \) on every variable that does not occur among the variables of \(\mathit{lits}\).
Let \(\mathit{lits}\) be a list of literals on \(n\) Boolean variables whose variables are pairwise distinct, let \(\rho \) be a restriction, and let \(g\) be a map sending each restriction to a decision tree. Let \(T\) be the sub-tree built for the term \(\mathit{lits}\) under \(\rho \) with continuation \(g\): reading the literals of \(\mathit{lits}\) in order, it branches on each literal whose variable is free under \(\rho \), skips the others, and evaluates \(g\) at each leaf. Write \(\mathit{lits}_\rho \) for the sublist of \(\mathit{lits}\), in order, consisting of exactly those literals whose variable is free under \(\rho \). Then for every index \(k\) smaller than both the length of \(\mathit{lits}_\rho \) and the length of the deepest root-to-leaf path of \(T\), the variable queried at position \(k\) of that deepest path equals the variable of the \(k\)-th literal of \(\mathit{lits}_\rho \).
Fix \(n\), and let \(\ell _1,\dots ,\ell _m\) be a list of literals over \(n\) variables, \(\rho \) a restriction on those variables, \(\mathrm{cont}\) a map sending each restriction to a decision tree, and \(x \in \{ \mathrm{true},\mathrm{false}\} ^n\) an assignment. Then evaluating the term sub-tree built from \(\ell _1,\dots ,\ell _m\), \(\rho \), and \(\mathrm{cont}\) at \(x\) gives the same Boolean value as evaluating \(\mathrm{cont}(\rho ^\ast )\) at \(x\), where \(\rho ^\ast \) is obtained from \(\rho \) by processing the literals in order: for each \(\ell _i\), if its variable is free in the current restriction, that variable is fixed to \(x\)’s value there, and otherwise the restriction is left unchanged.
Let \(\rho \) be a restriction on \(n\) variables, let \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\) be a point, and let \(\ell _1, \dots , \ell _m\) be a list of literals over these variables. Form a sequence of restrictions \(\rho _0 = \rho , \rho _1, \dots , \rho _m\) by processing the literals in order: for each \(k\), writing \(v_k\) for the variable of \(\ell _k\), if \(v_k\) is free in \(\rho _{k-1}\) then let \(\rho _k\) agree with \(\rho _{k-1}\) except that \(v_k\) is fixed to the value \(x(v_k)\), and otherwise let \(\rho _k = \rho _{k-1}\). Then the extension of \(\rho _m\) along \(x\) equals the extension of \(\rho \) along \(x\).
Let \(\rho \) be a restriction on \(n\) variables, let \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\) be a Boolean assignment, and let \(\ell _1, \dots , \ell _k\) be a list of literals on \(n\) variables. Process the literals from left to right, starting from \(\rho \): at each step, if the current literal’s variable is still free in the restriction obtained so far, fix that variable to the value \(x\) assigns it, and otherwise leave the restriction unchanged; write \(\rho '\) for the restriction produced after all \(k\) steps. If at least one literal in the list has its variable free in the original restriction \(\rho \), then \(\rho '\) has strictly fewer free variables than \(\rho \).
Let \(\rho : \mathrm{Fin}\, n \to \mathrm{Option}\, \mathrm{Bool}\) be a restriction on \(n\) variables, let \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\) be a Boolean assignment, and let \(\ell _1, \dots , \ell _m\) be literals on these variables. Process the literals from left to right, starting from \(\rho \): at each literal, if its variable is still free in the restriction accumulated so far, fix that variable to the value \(x\) assigns it, and otherwise change nothing. Then any variable \(v\) that already carries a fixed value under \(\rho \) (that is, \(\rho (v)\) leaves \(v\) set rather than free) still carries a fixed value in the final restriction; the fold never frees an already-fixed variable.
Fix \(n\) variables, let \(x : \mathrm{Fin}\, n \to \mathrm{Bool}\) be a Boolean assignment, let \(\rho \) be a restriction on these variables, and let \(\ell _1, \dots , \ell _m\) be a list of literals. Process the literals from left to right, maintaining a current restriction that begins as \(\rho \): at each literal \(\ell \), if its underlying variable \(\ell .\mathrm{var}\) is free in the current restriction, update that restriction by fixing \(\ell .\mathrm{var}\) to the value \(x(\ell .\mathrm{var})\), and otherwise leave the restriction unchanged; write \(\rho ^*\) for the restriction obtained after the whole list has been processed. Then for every literal \(\ell \) in the list whose variable is free in the original restriction \(\rho \), the value \(\rho ^*(\ell .\mathrm{var})\) is set, i.e. \(\ell .\mathrm{var}\) is not free in \(\rho ^*\).
Fix a restriction \(\rho \) on \(n\) variables and a continuation \(c\) assigning a decision tree to each restriction. For a list of literals \(L\), write \(T(L)\) for the term sub-tree built from \(L\) under \(\rho \) with continuation \(c\): it branches, in order, on each literal of \(L\) whose variable is free in \(\rho \), skipping the rest, and at every leaf invokes \(c\) on the restriction updated along that path. Let \(P\) and \(R\) be lists of literals, and suppose every literal in \(P\) has its variable already fixed by \(\rho \), that is, not among the free variables of \(\rho \). Then prepending \(P\) to \(R\) leaves the resulting tree unchanged:
where \(P \frown R\) denotes the concatenation of \(P\) followed by \(R\).
Let \(\rho \) be a restriction on \(n\) variables and \(g\) a continuation assigning a decision tree to each restriction on \(n\) variables. Fix a variable \(v\), a Boolean value \(b\), and a literal \(\ell \) whose variable is \(v\), and let \(\mathit{rest}\) be a further list of literals. Write \(\rho [v \mapsto b]\) for the restriction that agrees with \(\rho \) everywhere except at \(v\), where it is fixed to the value \(b\). Then the term sub-tree built from the list obtained by prepending \(\ell \) to \(\mathit{rest}\), under \(\rho [v \mapsto b]\) with continuation \(g\), coincides with the term sub-tree built from \(\mathit{rest}\) alone under the same restriction and the same continuation.
Fix a natural number \(w\) and let \(l\) be a finite list of pairs \((i, b)\) with \(i \in \mathbb {N}\) and \(b\) a Boolean, every first coordinate of which satisfies \(i \lt w\). Recasting \(l\) into a list of pairs in \(\mathrm{Fin}\, w \times \{ \text{true}, \text{false}\} \) — by reinterpreting each index \(i\) as the element of \(\mathrm{Fin}\, w\) it names, while leaving the Booleans untouched — yields a list of the same length as \(l\).
Fix a natural number \(w\), and let \(l\) be a finite list of pairs \((i,b)\) with \(i \in \mathbb {N}\) and \(b\) a Boolean, subject to the requirement that \(i \lt w\) for every pair in \(l\). Reinterpret each such pair as \((\bar\imath , b)\), where \(\bar\imath \) is the element of \(\mathrm{Fin}\, w\) whose underlying value is \(i\), obtaining a list over \(\mathrm{Fin}\, w \times \mathrm{Bool}\). Applying the map \((j,b) \mapsto (j\text{'s underlying natural number}, b)\) to each entry of that list yields \(l\).
Fix \(w \in \mathbb {N}\), and let \(l\) be a list of pairs \((i, b) \in \mathbb {N}\times \{ \mathrm{true}, \mathrm{false}\} \) each of whose first coordinates satisfies \(i \lt w\). If \(l\) is nonempty, then its cast into a list of pairs in \(\mathrm{Fin}\, w \times \{ \mathrm{true}, \mathrm{false}\} \) is again nonempty.
The inverse direction of SwitchingLemma.parseAux: converts a list of triples back into an aux list of \((\mathbb {N} \times \mathrm{Bool})\) entries, emitting a \((w, \mathrm{false})\) termination marker after each triple whose marker flag is true.
Fix a width \(w \in \mathbb {N}\), and let \(\Phi \) denote the triples-to-aux conversion, which turns a finite list of triples in \(\mathrm{Fin}(w) \times \mathrm{Bool} \times \mathrm{Bool}\) into an aux list of \((\mathbb {N}\times \mathrm{Bool})\) entries. Then \(\Phi \) carries concatenation to concatenation: for any two such lists \(T_1\) and \(T_2\),
Let \(w \in \mathbb {N}\) and let \(b\) be a nonempty list of pairs in \(\mathrm{Fin}\, w \times \mathrm{Bool}\). If one first marks the last entry of \(b\) (setting the marker flag to true on the final pair and false on all others) and then converts the resulting list of triples back into an aux list, the outcome is the list obtained from \(b\) by casting each pair \((i, d)\) to the entry \((i, d) \in \mathbb {N}\times \mathrm{Bool}\), followed by the single termination marker \((w, \mathrm{false})\).
Let \(t\) be a list over a type \(\alpha \), and tag each entry of \(t\) with its zero-based position, forming the list of pairs \((t_i, i)\). If a pair \((l, i)\) occurs among these tagged entries — equivalently, the entry of \(t\) at position \(i\) is \(l\) — then deleting the first \(i\) entries of \(t\) leaves a nonempty list whose first element is \(l\); that is, there is a list \(\mathit{rest}\) with \(t\) from position \(i\) onward equal to \(l\) prepended to \(\mathit{rest}\).
Let \(t = (\ell _0,\ell _1,\dots ,\ell _{m-1})\) be a term, that is a list of literals over \(\mathrm{Fin}\, n\), and let \(p\) be a Boolean predicate on literals. Annotate each literal with its position to obtain the list of pairs \((\ell _i, i)\), and retain those pairs whose literal satisfies \(p\); separately, retain the literals of \(t\) that satisfy \(p\). Then for every index \(k\) that is smaller than the length of both retained lists, the literal in the \(k\)-th retained pair equals the \(k\)-th retained literal of \(t\).
Let \(t\) be a finite list of elements of a type \(\alpha \), and tag each entry of \(t\) with its position to form the list of pairs \((t_i, i)\), whose second coordinates run through \(0, 1, \dots , |t| - 1\). Let \(p\) be any Boolean predicate on such pairs, and let \(w\) be a natural number with \(|t| \le w\). If a pair \((l, \mathit{idx})\) occurs in the sublist obtained by keeping exactly those tagged pairs on which \(p\) is true, then \(\mathit{idx} \lt w\).
Let \(t\) be a term over \(n\) variables, that is, a finite list of literals, and let \(p\) be a predicate on literals. Form from \(t\) the indexed list in which each literal is paired with its position in \(t\). Then the number of entries of this indexed list whose literal component satisfies \(p\) equals the number of literals of \(t\) that satisfy \(p\).
Let \(l\) be a list over a type \(\alpha \), let \(p\colon \alpha \to \{ \text{true},\text{false}\} \) be a predicate, and fix a starting index \(s\). Form the indexed list obtained by pairing each entry of \(l\) with its position, numbered consecutively from \(s\). If the first pair in this indexed list whose first component satisfies \(p\) is the pair \((x,\mathrm{idx})\), then \(x\) is the first entry of \(l\) satisfying \(p\).
Let \(f : (\mathrm{Fin}\, n \to \mathrm{Bool}) \to \mathrm{Bool}\) be a Boolean function, \(d\) a natural number, and \(\rho \) a restriction on the \(n\) variables. Then \(\rho \) is bad for the negation \(\lnot f\) at depth \(d\) if and only if \(\rho \) is bad for \(f\) at depth \(d\); that is, the restricted function of \(\lnot f\) under \(\rho \) has decision-tree depth greater than \(d\) exactly when the restricted function of \(f\) under \(\rho \) does.
Let \(f : (\mathrm{Fin}\, n \to \mathbb {F}_2) \to \mathbb {F}_2\) be a Boolean function and \(\rho \) a restriction on \(n\) variables. Then restricting the negation of \(f\) agrees, at every point \(x\), with the negation of the restriction of \(f\); that is, \((\lnot f)\) restricted by \(\rho \) equals the pointwise negation of \(f\) restricted by \(\rho \).
Fix \(n \ge 1\) and let \(\psi \) be a CNF formula on \(n\) Boolean variables whose width is at most \(w\), meaning every clause of \(\psi \) has at most \(w\) literals; assume moreover that no clause of \(\psi \) repeats a literal or uses a variable in two of its literals, and that \(5s \le n\). Call a restriction \(\rho \) that leaves exactly \(s\) of the \(n\) variables free bad if the restricted function \(x \mapsto \psi (\rho .\mathrm{extend}\, x)\) is computed by no DNF formula of width at most \(w\). Then the number \(B\) of bad restrictions satisfies
Let \(\psi \) be a CNF formula on \(n \ge 1\) Boolean variables whose width is at most \(w\), and suppose that within each clause no literal is repeated and any two literals on the same variable coincide, so that the literals of every clause lie on distinct variables. Let \(s\) and \(d\) be natural numbers with \(5s \le n\). For a restriction \(\rho \) leaving some variables free and fixing the rest, write \(\psi \! \restriction _\rho \) for the Boolean function \(x \mapsto \psi \bigl(\rho .\mathrm{extend}\, x\bigr)\) obtained by evaluating \(\psi \) with the fixed coordinates set by \(\rho \) and the free coordinates supplied by \(x\), and let \(\mathrm{depth}(\psi \! \restriction _\rho )\) denote its decision-tree depth, the least depth of a decision tree computing it. Then, counting restrictions that leave exactly \(s\) variables free and under which the restricted function still requires depth greater than \(d\),
Fix a prime \(p\) and a natural number \(n\), and let \(M \subseteq \{ 0,1,\dots ,n-1\} \) be a set of coordinates. For every \(s \in V\) and every \(v \in V_M\) (that is, every \(v\) supported on \(M\)),
In words: applying the restriction endomorphism \(r_M^V\), which zeroes out the coordinates outside \(M\), to the left argument of the symplectic form leaves the value unchanged whenever the right argument is supported on \(M\).
Let \(p\) be a prime, let \(n\) be a natural number, and let \(M \subseteq \{ 0,1,\dots ,n-1\} \) be a subset of coordinates. Write \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) for the symplectic space equipped with the form \(\omega \big((x,z),(x',z')\big) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\), and let \(V_M \subseteq V\) be the support submodule of those vectors whose coordinates outside \(M\) all vanish. If \(v \in V_M\) satisfies \(\omega (v,w) = 0\) for every \(w \in V_M\), then \(v = 0\).
Fix a prime \(p\) and an integer \(n\), and work in the space \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) equipped with the symplectic form \(\omega \). Let \(M \subseteq \{ 0, 1, \dots , n-1\} \) be a set of coordinates, and let \(v \in V_M\), so that every coordinate of \(v\) lying outside \(M\) vanishes. Then for every \(s \in V\),
where \(r_M(s)\) is the vector obtained from \(s\) by zeroing out all coordinates outside \(M\).
Let \(p\) be prime, \(\mathbb {F}_p\) the field with \(p\) elements, and let \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) carry the symplectic form \(\omega \big((x,z),(x',z')\big) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\). For a subset \(M \subseteq \{ 0,\dots ,n-1\} \), let \(r_M : V \to V\) be the endomorphism that keeps the coordinates indexed by \(M\) and sets all others to zero, and let \(V_M \subseteq V\) be the subspace of vectors whose support is contained in \(M\). Then for every \(s \in V\) and every \(v \in V_M\),
Let \(p\) be prime and work over the space \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) equipped with the symplectic form \(\omega (u,v) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\), where \(u = (x,z)\) and \(v = (x',z')\). Then for all \(x, y, z \in V\),
Fix a prime \(p\) and a natural number \(n\), and let \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\), where \(\mathbb {F}_p = \mathbb {Z}/p\mathbb {Z}\). Equip \(V\) with the symplectic form \(\omega (u,v) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\) for \(u = (x,z)\) and \(v = (x',z')\). Then \(\omega \) is additive in its second argument: for all \(x, y, z \in V\),
Let \(p\) be prime and work over the prime field \(\mathbb {F}_p\), with \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) carrying the symplectic form \(\omega \) given by \(\omega (u,v) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\) for \(u=(x,z)\) and \(v=(x',z')\). Then for every scalar \(c \in \mathbb {F}_p\) and all \(x, y \in V\),
where \(c\, x\) denotes the coordinatewise scalar multiplication of \(x\) by \(c\).
Let \(p\) be a prime, and equip \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) with the symplectic form \(\omega (u,v) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\), where \(u = (x,z)\) and \(v = (x',z')\). Then for every scalar \(c \in \mathbb {F}_p\) and all \(x, y \in V\),
where \(c\, y\) denotes componentwise scalar multiplication.
Fix a prime \(p\) and let \(\omega \) denote the symplectic form on \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\). For a subset \(M \subseteq \{ 0, \dots , n-1\} \) and any two vectors \(x, y\) in the support submodule \(V_M\), the restricted symplectic form evaluated at \((x, y)\) agrees with \(\omega (x, y)\) computed on \(x\) and \(y\) regarded as elements of the ambient space \(V\).
Fix a prime \(p\), a natural number \(n\), and a subset \(M \subseteq \{ 0,1,\dots ,n-1\} \), and let \(V_M\) be the associated support submodule of \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\). Then the restriction to \(V_M\) of the symplectic form \(\omega (u,v) = \sum _{i} (x_i z'_i - z_i x'_i)\), where \(u=(x,z)\) and \(v=(x',z')\), is reflexive: for all \(u, v \in V_M\), if \(\omega (u,v) = 0\) then \(\omega (v,u) = 0\).
Let \(p\) be a prime, and equip \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) with the symplectic form \(\omega \big((x,z),(x',z')\big) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\). For any subset \(M \subseteq \{ 0,\dots ,n-1\} \), let \(V_M\) be the support submodule consisting of those \((x,z) \in V\) with \(x_i = z_i = 0\) for every \(i \notin M\). Then the restriction of \(\omega \) to \(V_M\) is nondegenerate: if \(u \in V_M\) satisfies \(\omega (u,v) = 0\) for all \(v \in V_M\), then \(u = 0\).
Fix a prime \(p\) and an integer \(n\), and equip \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) with the bilinear form \(\omega (u,v) = \sum _{i=0}^{n-1}(x_i z'_i - z_i x'_i)\), where \(u = (x,z)\) and \(v = (x',z')\). Then \(\omega \) is antisymmetric: for all \(u, v \in V\),
Fix a prime \(p\) and an integer \(n\), and set \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\), equipped with the alternating bilinear form
Let \(B \colon V \times V \to \mathbb {F}_p\) be the bilinear form on \(V\) obtained by bundling \(\omega \). Then \(B\) agrees with \(\omega \) on every pair of arguments: for all \(u, v \in V\) one has \(B(u,v) = \omega (u,v)\).
Let \(p\) be a prime, \(n\) a natural number, and let \(\omega \) be the bilinear form on \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) given by \(\omega (u,v) = \sum _{i=1}^{n}(x_i z'_i - z_i x'_i)\) for \(u = (x,z)\) and \(v = (x',z')\). Then \(\omega \) is reflexive: for all \(u, v \in V\), if \(\omega (u,v) = 0\) then \(\omega (v,u) = 0\).
Fix a prime \(p\) and an integer \(n\), and let \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\) carry the bilinear form
over the prime field \(\mathbb {F}_p\). Then \(\omega \) is non-degenerate: the only vector \(u \in V\) with \(\omega (u,v) = 0\) for every \(v \in V\) is \(u = 0\).
Let \(k\), \(n\), \(d\) be natural numbers with \(k \gt 0\), \(n \ge 2\), and \(d \gt 0\), and let \(m\) be a natural number with \(m \le nk\). Then there exists a real constant \(C_2 \gt 0\) such that every real number \(v\) satisfying
also satisfies \(v \le C_2 \, n \, \sqrt{k} \, d \, (nk)^{1/8} (\log n)^{1/4}\).
Let \(f\) be a 3-CNF formula with \(m\) clauses, and let \(s\) be a set of vertices of the conflict graph of \(f\). Suppose that \(s\) has exactly \(m\) elements and forms a clique, meaning any two distinct vertices of \(s\) are adjacent. Then for every clause index \(i \in \mathrm{Fin}\, m\) there is exactly one vertex \(u \in s\) whose clause component equals \(i\).
Let \(f\) be a 3-CNF formula, let \(v\) be a vertex of the conflict graph of \(f\), and let \(i\) be a clause index of \(f\) such that the clause component of \(v\) equals \(i\). Then the literal named by \(v\) is one of the three literals \(l_1, l_2, l_3\) of the clause \(f[i]\).
Call two literals over a variable set *conflicting* when one is the positive and the other the negative occurrence of one and the same variable. For any two literals \(l_1\) and \(l_2\), \(l_1\) and \(l_2\) conflict if and only if \(l_2\) and \(l_1\) conflict; that is, the conflict relation is symmetric.
Let \(\alpha \) be a truth assignment on a set \(V\) of variables, and let \(\ell _1,\ell _2\) be literals over \(V\), each of the form \(\mathrm{pos}\, v\) (the variable \(v\)) or \(\mathrm{neg}\, v\) (its negation). If both \(\ell _1\) and \(\ell _2\) evaluate to true under \(\alpha \), then they do not conflict; that is, they are not the positive and negative occurrences of one and the same variable.
Let \(f\) be a 3-CNF formula consisting of \(m\) clauses, and let its conflict graph be the simple graph on the vertices \(\langle c, \ell \rangle \) (clause index \(c\), literal position \(\ell \)) in which two vertices are adjacent exactly when they belong to different clauses and their literals do not conflict. If \(f\) is satisfiable, then this conflict graph has an \(m\)-clique.
Let \(f\) be a 3-CNF formula with \(m\) clauses. Then \(f\) is satisfiable if and only if its conflict graph has a clique of size \(m\); that is, the graph whose vertices are the literal positions of \(f\), with two positions adjacent exactly when they lie in different clauses and their literals do not conflict, contains \(m\) pairwise-adjacent vertices precisely when some truth assignment satisfies every clause of \(f\).
The conflict graph of a 3-CNF formula \(f\) is the simple graph on vertex set \(\mathtt{CliqueVertex}\, f.\mathtt{length}\) in which two vertices \(u\) and \(v\) are adjacent if and only if they come from different clauses (\(u.c\_ idx \ne v.c\_ idx\)) and their respective literals do not conflict.
Let \(\alpha \) carry a linear order, and for \(x \in \alpha \) and a list \(\mathit{xs}\) over \(\alpha \) let \(\mathrm{insert}(x, \mathit{xs})\) denote the ordered insertion that scans \(\mathit{xs}\) from the front, spending one unit of time on each element compared, until it reaches the first entry not smaller than \(x\). Then the time cost incurred by \(\mathrm{insert}(x, \mathit{xs})\) is at most the length of \(\mathit{xs}\).
Let \(\alpha \) be a linearly ordered type and let \(xs\) be a finite list of elements of \(\alpha \). Running the timed insertion sort on \(xs\) charges at most \(T(\left\lvert xs\right\rvert )\) units of time, where \(\left\lvert xs\right\rvert \) is the length of \(xs\) and \(T\) is the recurrence given by \(T(0) = 0\) and \(T(n+1) = T(n) + n\).
For a cost type \(T\) (typically \(\mathbb {N}\)) and a value type \(\alpha \), a timed computation is a pair consisting of a return value \(\mathrm{ret} \in \alpha \) together with an accumulated time cost \(\mathrm{time} \in T\). This is the carrier of the time-complexity monad: correctness properties are stated about the return value, and complexity bounds about the recorded cost.
Given \(m : \texttt{TimeM T α}\) and \(f : \alpha \to \texttt{TimeM T β}\), the bind of \(m\) and \(f\) runs \(f\) on the value returned by \(m\) and returns \(\langle (f\, m.\mathrm{ret}).\mathrm{ret},\; m.\mathrm{time} + (f\, m.\mathrm{ret}).\mathrm{time}\rangle \): the two time costs are added.
Consider computations that carry both a return value and an accumulated time cost drawn from a type \(T\) equipped with an addition, and let the monadic bind of such a computation \(m\) with a continuation \(f\) run \(m\), pass its return value to \(f\), and add the two costs. Then the return value of \(m \mathbin {\gt \! \! \gt \! \! =} f\) equals the return value of \(f\) applied to the return value of \(m\).
A value of type \(\mathrm{TimeM}\ T\ \alpha \) is a timed computation, consisting of a return value \(\mathrm{ret} \in \alpha \) together with an accumulated time cost in \(T\), and \(\mathrm{TimeM}\ T\) carries the functor structure that acts on the return value. For any function \(f : \alpha \to \beta \) and any timed computation \(x\) of type \(\mathrm{TimeM}\ T\ \alpha \), the return value of the mapped computation obtained by applying \(f\) through this functor to \(x\) equals \(f(\mathrm{ret}(x))\).
A *timed computation* \(\mathrm{TimeM}\, T\, \alpha \) records a return value in \(\alpha \) together with an accumulated time cost in \(T\). Suppose \(T\) carries a distinguished zero element, and let \(\mathrm{pure}\, a\) denote the timed computation that returns \(a\) at zero cost. Then for every \(a \in \alpha \), the return value of \(\mathrm{pure}\, a\) is \(a\).
Regard a value of type \(\mathrm{TimeM}\, T\, \alpha \) as a timed computation, carrying both a return value in \(\alpha \) and an accumulated time cost in \(T\), and suppose \(T\) is equipped with an addition, so that such computations may be sequenced. Let \(f\) be a timed computation whose return value is a function \(\alpha \to \beta \), and let \(x\) be a thunk, that is, a function on the one-point type producing a timed computation with return value in \(\alpha \). Then the return value of the computation obtained by sequencing \(f\) with \(x\) equals the return value of \(f\) applied to the return value of the computation \(x(\, )\) obtained by forcing the thunk.
A value of type \(\mathrm{TimeM}\, T\, \alpha \) is a computation carrying both a return value in \(\alpha \) and an accumulated cost in \(T\); write \(\mathrm{ret}\) for the return-value component. Suppose \(T\) is equipped with an addition, let \(x\) be such a computation of type \(\mathrm{TimeM}\, T\, \alpha \), and let \(y\) be a (thunked) computation of type \(\mathrm{TimeM}\, T\, \beta \). Then the left-sequencing \(x \mathbin {\lt {*}} y\), which runs both computations and adds their costs but keeps only the value produced by \(x\), has return value \(\mathrm{ret}(x \mathbin {\lt {*}} y) = \mathrm{ret}(x)\).
Model a computation that tracks its running time as a pair consisting of a return value and an accumulated time cost drawn from a type \(T\) equipped with an addition. Given such a computation \(x\), with return value in a type \(\alpha \), and a thunk \(y\) producing a second such computation with return value in a type \(\beta \), the right-sequencing of \(x\) and \(y\) — which runs \(x\), then \(y()\), combines their time costs, and discards the value of \(x\) — has return value equal to that of \(y()\).
Model a computation as a pair consisting of a return value in a type \(\alpha \) together with an accumulated time cost drawn from a type \(T\) equipped with an addition operation; write \(\mathrm{ret}(\cdot )\) for the returned value and \(\mathrm{time}(\cdot )\) for the cost. Given such a computation \(m\) and a continuation \(f\) that assigns to each value of type \(\alpha \) a further computation returning a value of type \(\beta \), the sequential composition \(m \gg \! = f\) (run \(m\), then run \(f\) on its result) has cost equal to the cost of \(m\) plus the cost of running \(f\) on the value returned by \(m\):
A timed computation over a cost type \(T\) is a pair consisting of a return value in \(\alpha \) and an accumulated time cost in \(T\). For any function \(f\colon \alpha \to \beta \) and any timed computation \(x\), applying \(f\) to the return value of \(x\) (leaving the value’s type as \(\beta \)) yields a timed computation whose time cost equals that of \(x\).
A cost-annotated computation over a type \(T\) of costs and a value type \(\alpha \) consists of a return value in \(\alpha \) together with an accumulated time cost in \(T\). Suppose \(T\) carries a distinguished zero element \(0\). Then for every value \(a \in \alpha \), the trivial computation that returns \(a\) (the \(\mathrm{pure}\) computation) has time cost \(0\).
Let \(T\) be a type equipped with an addition, and regard a value of \(\mathrm{TimeM}\, T\, \alpha \) as a computation that returns a result in \(\alpha \) together with an accumulated time cost \(\mathrm{time}(\cdot ) \in T\). Given a computation \(f\) of a function \(\alpha \to \beta \) and a thunked computation \(x\) of an argument in \(\alpha \), the applicative sequencing \(f \mathbin {\langle *\rangle } x\), which applies the computed function to the computed argument, has time cost \(\mathrm{time}(f) + \mathrm{time}(x())\), the sum of the cost of \(f\) and the cost of forcing and running \(x\).
Fix a type \(T\) carrying an addition, and regard a value of type \(\mathrm{TimeM}\, T\, \alpha \) as a computation that records both a return value in \(\alpha \) and an accumulated time cost \(\mathrm{time}\in T\). Given such a computation \(x\) and a thunked computation \(y\colon \mathrm{Unit}\to \mathrm{TimeM}\, T\, \beta \), the left-sequencing \(x \mathbin {\lt *} y\) — which runs both but keeps the return value of \(x\) — has time cost \(\mathrm{time}(x) + \mathrm{time}\bigl(y()\bigr)\), the sum of the two individual costs.
A *timed computation* over an additive cost type \(T\) records a return value together with an accumulated cost in \(T\); write \(\mathrm{time}(c) \in T\) for the cost of such a computation \(c\). Let \(x\) be a timed computation and let \(y\) be a thunk that, when forced, produces a timed computation. The right-sequencing \(x \mathbin {*\gt } y\), which performs \(x\) and then \(y()\) and returns the value of \(y()\), has time cost
Let \(A\in \mathbb {R}^{k\times d}\) be a \(k\times d\) real matrix, regarded as the linear map from the Euclidean space \(\mathbb {R}^{d}\) to \(\mathbb {R}^{k}\) that it induces, and let \(x\in \mathbb {R}^{d}\). Then for each index \(i\in \{ 1,\dots ,k\} \), the \(i\)-th coordinate of the image \(Ax\) is given by
Let \(D\) be a tournament on a finite vertex set \(V\) (an orientation of the complete graph on \(V\)), and let \(A\) be its \(0/1\) adjacency matrix. Then \(A\) is primitive — that is, some power \(A^{k}\) has all entries positive — if and only if \(D\) is diconnected (every vertex is reachable from every other) and \(V\) has at least four vertices.
Let \(D\) be a diconnected tournament on a finite vertex set \(V\) with \(\nu = \left\lvert V\right\rvert \ge 5\), and let \(A\) be its \(\nu \times \nu \) adjacency matrix over \(\mathbb {N}\), so that \(A_{uv} = 1\) when there is an arc from \(u\) to \(v\) and \(A_{uv} = 0\) otherwise. Writing \(d\) for the directed diameter of \(D\), every entry of the matrix power \(A^{\, d+3}\) is strictly positive; that is, \(\bigl(A^{\, d+3}\bigr)_{ij} \gt 0\) for all \(i, j \in V\).
Let \(D\) be a tournament on a finite, nonempty vertex set \(V\). Then either \(D\) is diconnected, or there is an arc \(x \to y\) of \(D\) whose reversal makes it so: deleting the arc \(x \to y\) and inserting the arc \(y \to x\), while leaving every other arc unchanged, yields a diconnected digraph.
Let \(D\) be a tournament on a finite, nonempty vertex set \(V\); that is, for each pair of distinct vertices exactly one of the two directed edges between them is present, and no vertex has an edge to itself. Write \(u \to v\) when the edge from \(u\) to \(v\) is present. Then \(D\) has a king: a vertex \(u\) such that every other vertex \(v\) is reached from \(u\) by a directed path of length at most two — that is, either \(u \to v\), or there is a vertex \(w\) with both \(u \to w\) and \(w \to v\).
Let \(D\) be a tournament on a finite vertex set \(V\) — an orientation of the complete graph on \(V\). Let \(v_1, v_2, \dots , v_k\) be a directed path in \(D\): a nonempty list of distinct vertices such that \(v_j \to v_{j+1}\) is an arc of \(D\) for every \(1 \le j \lt k\). Suppose this path cannot be extended at either end, meaning that for every vertex \(w\) not among \(v_1, \dots , v_k\) there is no arc from \(w\) to \(v_1\) and no arc from \(v_k\) to \(w\). Then for every vertex \(v\) not on the path there is an index \(i\) with \(1 \le i \lt k\) such that \(v_i \to v\) and \(v \to v_{i+1}\) are both arcs of \(D\); that is, \(v\) can be inserted between two consecutive vertices of the path.
An action of a \(k\)-tape Turing machine over a symbol alphabet and a set of states records everything the machine does in one step: an attempted movement of the input head (left, stay, or right, encoded as a sign in \(\{ -1,0,1\} \)); for each of the \(k\) work tapes, optionally a value to write into the scanned cell (either a symbol or the blank) together with a head movement; an optional symbol to emit on the output tape; and an optional successor state, where the absence of a successor state means the machine halts.
The result of carrying out an action on a configuration: the internal state becomes the action’s (optional) successor state; the input head is moved by the action’s sign using the clamped move, so it never leaves the input region and its two flanking blank cells; on each work tape, the scanned cell is overwritten with the action’s value for that tape when one is prescribed (and left unchanged otherwise) and the head then moves by the action’s sign for that tape; and the action’s emitted symbol, if any, is appended to the output tape. This is the part of a machine step that does not depend on how the action was chosen.
Let \(a\) be a one-step action of a multi-tape Turing machine with \(k\) work tapes, let \(c\) be a configuration of such a machine, and let \(i\) be one of the \(k\) work-tape indices. After applying \(a\) to \(c\), the content of the \(i\)-th work tape equals the old content updated at the cell under the \(i\)-th work head, where the new cell value is the symbol that \(a\) proposes to write there and, when \(a\) declines to write on that tape, the symbol previously scanned: an explicit blank write erases the cell, while no write at all copies the scanned symbol back unchanged. No finiteness or liveness hypothesis is required.
The extension of a one-step action on \(k\) work tapes to an action on \(k+1\) work tapes: on the first \(k\) tapes it writes and moves exactly as before, and on the extra last tape it writes nothing and does not move the head. The input-head move, the output emission, and the next state are unchanged.
Given a function \(f\) from one state set to another, the one-step action of a Turing machine with \(k\) work tapes obtained from an action \(a\) by applying \(f\) to its optional successor state. The input-head movement, the write-and-move instructions for the \(k\) work tapes, and the optional output symbol are unchanged, and a halting action (one with no successor state) remains halting.
The serialization of a single transition record of a one-work-tape binary machine with \(n + 1\) states: the concatenation of the two-bit input-head move, the two-bit optional write on the work tape, the two-bit work-head move, the two-bit optional output emission, and the self-delimiting serialization of the optional successor state.
Let \(c\) be a program in Mathlib’s partial-recursive language and \(x\) a Boolean input string. For every position \(j \le \left\lvert x\right\rvert \): from the bridge configuration of the bridging machine of \(c\) in the backward-reading state, with input head at position \(j\) and the main stack holding the stack word of the suffix of \(x\) from position \(j\) onward (other stacks empty, empty output), the machine reaches the configuration with the input head at the left endmarker whose main stack holds the stack word of all of \(x\) and whose control is the execution state at the translated initial statement of \(c\)’s compilation, with empty register.
Let \(c\) be a program in Mathlib’s partial-recursive language and \(f\) a function on Boolean strings such that evaluating \(c\) on the singleton list holding the sentinel numeric code of each string \(x\) yields exactly the singleton list holding the code of \(f(x)\). Then there exist a finite Turing machine over the two-letter Boolean alphabet and a time bound \(T \colon \mathbb {N}\to \mathbb {N}\) such that the machine computes \(f\), halting within \(T(\left\lvert x\right\rvert )\) steps on every input \(x\).
Let \(c\) be a program in Mathlib’s partial-recursive language and \(f\) a total function on Boolean strings such that, for every string \(x\), evaluating \(c\) on the singleton list containing the sentinel numeric code of \(x\) yields exactly the singleton list containing the sentinel numeric code of \(f(x)\). Then for every \(x\) there exists a time \(t\) such that the bridging machine of \(c\), run on the embedded input \(x\), halts within \(t\) steps with the embedded string \(f(x)\) on its output tape.
Suppose that in a bridge configuration of the bridging machine the main stack holds the stack word of a Boolean string \(xs\) and the control is the emitting state with an optional pending bit. Then the machine reaches the halted configuration whose main stack retains only the list terminator (the other stacks unchanged) and whose output is extended by the pending bit, if present, followed by the bits of \(xs\), all embedded into the bridge alphabet; the trailing sentinel bit of the stack word is consumed without being emitted.
Suppose the bridging machine is in the emitting state with an optional pending bit, and the main stack begins with the bit symbol embedding a Boolean \(b\). Then a single step pops that symbol, appends the pending bit (if present, embedded into the bridge alphabet) to the output, and continues in the emitting state with pending bit \(b\); the input position and the other stacks are unchanged.
Let \(c\) be a program in Mathlib’s partial-recursive language and \(x\) a Boolean input string. For every \(j \le \left\lvert x\right\rvert \), running the bridging machine of \(c\) for exactly \(j\) steps from its initial configuration on the embedded input \(x\) yields the bridge configuration in the scanning state with the input head at position \(j + 1\) (positions are shifted by one, position \(0\) being the left endmarker), all stacks empty, and empty output.
From the bridge configuration of the bridging machine in the scanning state at the right endmarker of the embedded Boolean input \(x\) (all stacks empty, empty output), exactly three steps lead to the backward-reading state with the input head at position \(\left\lvert x\right\rvert \) and the main stack holding just the sentinel true-bit symbol on top of the list terminator, the other stacks empty and the output empty.
Let \(c\) be a program in Mathlib’s partial-recursive language, and let \(a, b\) be configurations of the compiled stack machine such that \(b\) is reachable from \(a\) and the label of \(a\) lies in the label support of \(c\)’s compilation with the halt label adjoined. Then the label of \(b\) also lies in that support, and, for every fixed input head position and output string, the bridging machine of \(c\) reaches, from the bridge configuration associated with \(a\) (the optional execution state for \(a\)’s translated label and register value, over \(a\)’s stacks), the bridge configuration associated with \(b\).
From its initial configuration on the embedded Boolean input \(x\), the bridging machine of a partial-recursive program \(c\) reaches the bridge configuration at input position \(0\) whose main stack holds the full stack word of \(x\) (other stacks empty, empty output), in the execution state at the translated initial statement of \(c\)’s compilation, with empty register.
Let \(c\) be a program in Mathlib’s partial-recursive language, \(q\) a stack-machine statement in the finite statement support of \(c\)’s compilation, \(v\) an optional register value, and \(st\) a stack assignment. Executing \(q\) on \(v\) and \(st\) in the stack-machine semantics produces a next optional label, register value, and stacks. From the bridge configuration in the execution state for \(q\) and \(v\) over \(st\), the bridging machine of \(c\) reaches, in finitely many steps, the bridge configuration in the optional execution state determined by the translation of that next label and by that register value, over those resulting stacks, with the input position and output unchanged.
One step of the bridging machine from a bridge configuration in a live control state is computed by applying the action that its transition table prescribes for that state, the symbol under the input head, and the symbols under the four work-tape heads.
The bridge configuration determined by an optional control state, an input head position, an assignment \(st\) of stacks to the four stack names, and an output string: each stack \(st(k)\) is laid out on the corresponding work tape in the standard representation, with that tape’s head at the top position \(-\left\lvert st(k)\right\rvert \).
Suppose that in a bridge configuration the stack assigned to a name \(k\) is nonempty, with top symbol \(a\) above a tail \(xs\). Then the one-tape action on the tape of \(k\) that blanks the current cell and moves that head one cell to the right, passing to a control state \(q'\), transforms the configuration into the bridge configuration with control state \(q'\) whose assignment replaces the stack at \(k\) by \(xs\), the input position and output unchanged.
The one-tape action on the tape of a stack name \(k\) that blanks the current cell and moves the head right precisely when the stack at \(k\) is nonempty transforms a bridge configuration into the bridge configuration whose stack at \(k\) is replaced by its tail (an empty stack remaining empty), with the input position and output unchanged.
Suppose that in a bridge configuration the stack at a name \(k\) is nonempty, with top symbol \(a\) above a tail \(xs\). The one-tape pop action on the tape of \(k\) (blank the current cell, move right), extended with the emission of an optional output symbol \(e\), produces the bridge configuration whose stack at \(k\) is \(xs\) and whose output is extended by \(e\), the input position unchanged.
Pushing a symbol \(a\) onto the stack named \(k\) is realized by two consecutive one-tape actions on the tape of \(k\): first move that head one cell to the left (writing nothing), then write the embedded symbol \(a\) without moving. Applied to a bridge configuration, this yields the bridge configuration whose assignment replaces the stack at \(k\) by \(a\) prepended to it, with the specified final control state and the input position and output unchanged.
Performing the two-action push of a symbol \(a\) onto the stack named \(k\), where the second action additionally moves the input head in a direction \(d\), transforms a bridge configuration into the bridge configuration with \(a\) prepended to the stack at \(k\) and the input head moved by \(d\), the output unchanged.
The optional execution control state determined by an optional stack-machine statement \(q\) and an optional register value \(v\): the execution state carrying \(q\) and \(v\) when \(q\) belongs to the finite statement support of the compiled program, and no state (immediate halt) otherwise.
If an optional statement \(q\) belongs to the finite statement support of the compiled program, then the optional execution control state determined by \(q\) and a register value \(v\) is precisely the execution state carrying \(q\) and \(v\).
The sentinel numeric code of a Boolean string: the natural number whose binary expansion, least-significant bit first, consists of the string’s bits followed by one final true bit. It is obtained by folding the bit-append operation \((b, n) \mapsto 2n + b\) over the string, starting from \(1\).
A one-tape action of the bridging machine: on a selected work tape, optionally write a symbol and move that head in a given direction, leaving the input head and all other work tapes untouched, emitting no output, and passing control to a given optional successor state.
The work-tape contents representing a stack: a stack \(xs\) over the compiled stack machine’s alphabet occupies the cells \(-\left\lvert xs\right\rvert , \dots , -1\) of a work tape, with the top of the stack at cell \(-\left\lvert xs\right\rvert \), each symbol embedded into the bridge alphabet, and every other cell blank.
The control states of the bridging machine over a type \(Q\) of source statements: a forward input-scanning state; two seeding states that lay down the list terminator and the sentinel bit; a backward-reading state and a state that pushes one read input bit; execution and push states carrying a source statement from \(Q\) together with the optional register contents; and an emitting state carrying one optional pending output bit.
The finite statement support of a compiled program: for a program \(c\) in Mathlib’s partial-recursive language, the finite set of optional stack-machine statements arising from the labels in the support of \(c\)’s compilation, together with the halt marker.
For a program \(c\) in Mathlib’s partial-recursive language, the bridging machine of \(c\): a finite Turing machine over the bridge alphabet with four work tapes, one per stack of the compiled stack machine. Each stack occupies the negative cells of its tape ending at \(-1\), with the head on the stack top and a blank head at zero for an empty stack. The machine scans its input to the right end, seeds the main tape with the list terminator and a sentinel true bit, reads the input backwards while pushing each bit onto the main tape, executes the compiled program’s statements over the finite statement support, and finally pops the main stack while emitting the result bits to the output tape.
Let \(S\) and \(H\) be sets of states, \(\iota : S \to H\), \(r \in H\), and \(p\), \(o\) Boolean words. For every action \(a\) and every configuration \(c\) of a Boolean Turing machine with \(k\) work tapes and state set \(S\) on an input \(x\), applying the capture transform of \(a\) to the capture image of \(c\) gives the capture image of the configuration obtained by applying \(a\) to \(c\). Here the capture transform of \(a\) performs \(a\) on the input and the first \(k\) of \(k+1\) work tapes, writes \(a\)’s emitted bit (if any) on the last tape and moves that head right, emits nothing, and maps the successor through \(\iota \) with a halt replaced by \(r\); and the capture image of \(c\) has state \(\iota (q)\) (or \(r\) if \(c\) is halted), the input head and work tapes of \(c\) on the first \(k\) tapes, the word \(p\) followed by the output of \(c\) stored from cell \(0\) of the last tape with that head just past its end, and output \(o\).
Let \(M\) be a Turing machine over the Boolean alphabet with \(k\) work tapes and state set \(S\), let \(N\) be one with \(k+1\) work tapes and state set \(H\), and let \(\iota : S \to H\) and \(r \in H\). Suppose that for every \(s \in S\), every (possibly blank) input symbol \(\sigma \) and every tuple \((w_0, \dots , w_k)\) of work-tape symbols, the transition of \(N\) at \((\iota (s), \sigma , w_0, \dots , w_k)\) is the capture transform of the transition of \(M\) at \((s, \sigma , w_0, \dots , w_{k-1})\): \(M\)’s action on the input and first \(k\) tapes, its emitted bit (if any) written on the last tape followed by a right move, no output, and a halt replaced by entering \(r\). Let \(p\), \(o\) be Boolean words, \(c_0\) a configuration of \(M\) on an input \(x\), and \(t \in \mathbb {N}\) such that \(M\) run from \(c_0\) is not halted at any time \(t' \lt t\); then running \(N\) for \(t\) steps from the capture image of \(c_0\) yields the capture image of the configuration \(M\) reaches after \(t\) steps from \(c_0\), where the capture image of a configuration \(c\) has state \(\iota (q)\) (or \(r\) if \(c\) is halted), the input head and work tapes of \(c\) on the first \(k\) tapes, the word \(p\) followed by the output of \(c\) stored from cell \(0\) of the last tape with that head just past its end, and output \(o\).
Let \(S\) and \(H\) be sets of control states, \(\iota : S \to H\) a map and \(r \in H\) a designated return state. Given an action \(a\) of a Boolean Turing machine with \(k\) work tapes and state set \(S\), its capture transform is the action of a machine with \(k+1\) work tapes and state set \(H\) that moves the input head as \(a\) does and performs \(a\)’s write and head move on each of the first \(k\) work tapes; on the last work tape it writes the bit \(b\) and moves right if \(a\) emits \(b\), and neither writes nor moves if \(a\) emits nothing; it emits nothing on the output tape; and its successor state is \(\iota (q')\) if \(a\) has successor state \(q'\), and \(r\) if \(a\) halts.
Let \(S\) and \(H\) be sets of control states, \(\iota : S \to H\) a map, \(r \in H\) a return state, and \(p\), \(o\) Boolean words. For a configuration \(c\) of a Boolean Turing machine with \(k\) work tapes and state set \(S\) on an input \(x\), with output word \(u\), its capture image is the configuration on \(x\) of a machine with \(k+1\) work tapes and state set \(H\) whose control state is \(\iota (q)\) if \(c\) is in state \(q\) and \(r\) if \(c\) is halted; whose input head and first \(k\) work tapes (contents and head positions) are those of \(c\); whose last work tape holds the word \(pu\) in cells \(0, \dots , |pu|-1\), blank elsewhere, with its head at cell \(|pu|\); and whose output tape holds \(o\).
A configuration of a \(k\)-tape Turing machine, relative to a fixed input string of length \(n\), consists of: an optional internal state (absence of a state denoting the halting state); the position of the input head, one of the \(n + 2\) positions obtained by flanking the input with one blank cell on each side (so the head position is shifted by one relative to the input indices); the contents of the \(k\) work tapes, each a function from \(\mathbb {Z}\) to optional symbols (blank cells carrying no symbol); the integer position of the head on each work tape; and the contents of the write-only output tape, a finite string of symbols.
The embedding of a configuration of a machine with \(k\) work tapes and state set \(Q\) into a configuration of a machine with \(k+1\) work tapes over the extended state set \(Q \sqcup \{ 0,1,2\} \): the state (if the machine has not halted) is renamed along the left inclusion, the input head, the first \(k\) work tapes with their head positions, and the output are unchanged, and the extra last work tape is blank with its head at the origin.
Let \(a\) be a one-step action of a machine with \(k\) work tapes and state set \(Q\), and let \(c\) be a configuration of such a machine. Rename the states of \(a\) along the left inclusion \(Q \to Q \sqcup \{ 0,1,2\} \) and extend the result by an inert \((k+1)\)-st tape. Applying this extended, renamed action to the embedding of \(c\) into the \((k+1)\)-tape oracle model yields exactly the embedding of the configuration obtained by applying \(a\) to \(c\).
For any state \(q_0\) of a machine with \(k\) work tapes and any input word \(x\), the embedding into the \((k+1)\)-tape oracle model of the initial configuration in state \(q_0\) on input \(x\) is the initial configuration on input \(x\) whose state is the image of \(q_0\) under the left inclusion into the extended state set \(Q \sqcup \{ 0,1,2\} \).
For any configuration \(c\) of a machine with \(k\) work tapes and any of the \(k\) original tape indices \(i\), the symbol scanned on tape \(i\) in the embedding of \(c\) into the \((k+1)\)-tape oracle model equals the symbol scanned on tape \(i\) in \(c\).
The configuration in which a machine starts on a given input, for a starting state \(q_0\): the internal state is \(q_0\), the input head is at position \(1\) (scanning the first input symbol), every cell of every work tape is blank, every work tape head is at cell \(0\), and the output tape is empty.
The symbol currently scanned by the input head of a configuration: if the head sits on either boundary blank cell (position \(0\) or position \(n + 1\) for an input of length \(n\)), no symbol is read; otherwise, when the head is at position \(p\) with \(1 \le p \le n\), the symbol read is the \(p\)-th symbol of the input string.
Given a function \(f\) from one state set to another, the configuration of a Turing machine with \(k\) work tapes on a fixed input word obtained from a configuration \(c\) by applying \(f\) to its optional state component. The input-head position, the work tapes and their head positions, and the output are preserved, and a halted configuration (one with no state) stays halted.
Let \(f\) be an arbitrary function from one state set to another (no injectivity is required), let \(a\) be a one-step action of a Turing machine with \(k\) work tapes, and let \(c\) be a configuration on a fixed input word. Then applying the action \(a\) with states renamed along \(f\) to the configuration \(c\) with state renamed along \(f\) yields the same configuration as first applying \(a\) to \(c\) and then renaming the state of the result along \(f\).
Let \(k \ge 0\), let \(x\) be a Boolean input word, let \(q\) be a state, and let \(w = (w_0, \dots , w_{k-1})\) be a family of Boolean words indexed by the \(k\) work tapes. The seam configuration \(\mathrm{seam}(q, w)\) of a \(k\)-work-tape machine over the Boolean alphabet on input \(x\) is the configuration whose control is in the (non-halted) state \(q\), whose input head sits on the first input cell (position \(1\)), whose work tape \(i\) holds the word \(w_i\) in cells \(0, \dots , \left\lvert w_i\right\rvert - 1\) and is blank at every other integer cell, whose work-tape heads are all at cell \(0\), and whose output tape is empty.
Let \(k \ge 0\) be a number of work tapes, \(x\) a Boolean input word, \(q\) a state, \(w = (w_0, \dots , w_{k-1})\) a family of Boolean words, and \(i \lt k\) a work-tape index. In the seam configuration \(\mathrm{seam}(q, w)\) of a \(k\)-work-tape machine over the Boolean alphabet on input \(x\) (state \(q\), input head on the first input cell, work tape \(j\) holding \(w_j\) from cell \(0\), all work heads at cell \(0\), empty output), work tape \(i\) is exactly the buffer tape of \(w_i\): for every integer cell \(z\), its content is the \(z\)-th letter of \(w_i\) (counting from \(0\)) when \(0 \le z \lt \left\lvert w_i\right\rvert \), and blank otherwise.
The natural number denoted by a bit string read in least-significant-bit-first binary, so the string \(b_0 b_1 \cdots b_{k-1}\) denotes \(\sum _{j\lt k} b_j 2^j\). The function is total: it also assigns a value to noncanonical strings (for instance those with trailing zeros); canonicity is checked separately by the parser.
The canonical serialization denoted by a bit string \(x\): if the erased suffix scanner accepts \(x\) with unconsumed suffix \(t\), the result is the prefix of \(x\) of length \(|x| - |t|\) (the consumed part); if the scanner rejects \(x\), the result is the serialization of the fixed fallback machine.
There exist a finite Turing machine over the two-letter Boolean alphabet and a time bound \(T \colon \mathbb {N}\to \mathbb {N}\) such that on every Boolean string the machine halts within \(T\) of the input’s length, outputting the fixed serialization of the one-work-tape binary machine that the string decodes to (malformed strings decoding to the trivial one-state halting machine).
For an optional parse result \(r\) — either failure or a pair of a value and a suffix — and a partial continuation \(f\) on suffixes: binding \(r\) through the function applying \(f\) to the suffix component equals first discarding the value component of \(r\) and then binding through \(f\).
For every state-count parameter \(n\) and bit string \(x\): discarding the parsed record from the transition-record reader with parameter \(n\) yields exactly the transition-record skipper with parameter \(n\), which validates the four dictionary fields and the successor field without constructing the record.
For every bound \(n\) and bit string \(x\): discarding the parsed index from the result of the range-checked unary reader (keeping only the suffix, or failure) yields exactly the result of the unary index skipper with bound \(n\) on \(x\).
For every bit string \(x\): discarding the parsed movement from the head-movement reader yields exactly the two-bit skipper whose validity predicate accepts a pair \((a, b)\) precisely when \(a\) is \(1\) or \(b\) is \(0\) (that is, it rejects only the pair \(01\)).
For every field reader with values in a type \(A\), every length \(n\), and every bit string \(x\): discarding the parsed vector from the length-\(n\) vector reader equals iterating the value-discarding version of the field reader \(n\) times.
If the doubled-bit header parser succeeds on a bit string \(x\), returning a data word \(a\) and a suffix \(r\), then \(x\) is exactly the doubled-bit encoding of \(a\) — each bit of \(a\) repeated twice, then the separator \(01\) — followed by \(r\). This includes the case of an empty data word.
The partial parser for the fixed serialization of coded machines. On a bit string it: reads a doubled-bit header (each data bit repeated twice, terminated by the separator \(01\)) and requires the recovered word to be the canonical little-endian binary representation of its value \(n\); rejects unless the remaining input has length at least \(81(n+1)\); reads a unary initial state smaller than \(n + 1\); reads the full transition table, namely for each of the \(n + 1\) states a three-by-three grid of transition records indexed by the input and work symbols in blank/\(0\)/\(1\) order; and finally requires every remaining bit to be \(1\). On success it returns the coded machine with \(n + 1\) states assembled from these data.
For every bit string \(x\), the result of the machine parser on \(x\) equals the result of the suffix-retaining machine parser on \(x\) with the suffix component discarded: both fail together, and on success they return the same machine.
A variant of the machine parser following exactly the same grammar — doubled-bit canonical state count, minimum-length guard, unary initial state, full transition table, and an all-\(1\) remainder — but returning the decoded coded machine together with the unconsumed all-\(1\) suffix, rather than the machine alone.
Every primitive recursive function \(f\) on Boolean strings is computed by an actual machine: there exist a finite Turing machine over the two-letter Boolean alphabet and a time bound \(T \colon \mathbb {N}\to \mathbb {N}\) such that on every input \(x\) the machine halts within \(T(\left\lvert x\right\rvert )\) steps with \(f(x)\) on its output tape. No growth or effectivity condition is imposed on \(T\).
The canonizing string function is primitive recursive. This function maps a Boolean string \(xs\) to the prefix consumed by the erased machine-code parser when that parser accepts (namely the first \(\left\lvert xs\right\rvert - s\) bits, where \(s\) is the length of the leftover all-true suffix), and to the fixed serialization of the trivial one-state halting machine when the parser rejects.
The aligned pair parser is primitive recursive. This parser reads a Boolean string in aligned two-bit blocks: the blocks \(00\) and \(11\) contribute the data bits false and true, the first aligned block \(01\) acts as separator, and any other block (or a missing separator) is rejected; on success it returns the decoded first component together with the suffix following the separator.
The two-argument function taking a Boolean string \(xs\) and a natural number \(n\) to the prefix of \(xs\) of length \(\left\lvert xs\right\rvert - n\) (truncated subtraction, so the empty string when \(n \ge \left\lvert xs\right\rvert \)) is primitive recursive.
Let \(A\) be a type with a primitive recursive coding, let \(r\) assign to each \(a \in A\) a partial reader on Boolean strings, primitive recursively jointly in \(a\) and the string, and let \(\mathrm{count} \colon A \to \mathbb {N}\) be primitive recursive. Then the function taking \(a\) and a string to the result of applying the reader \(r(a)\) sequentially \(\mathrm{count}(a)\) times is primitive recursive jointly in \(a\) and the string.
The erased machine-code parser is primitive recursive. On a Boolean string this parser recovers a state count \(n\) from the aligned doubled-bit region, rejects unless those bits are exactly the canonical binary expansion of \(n\), rejects when fewer than \(81(n+1)\) bits remain, skips a unary initial state below \(n + 1\) and then \(9(n+1)\) five-field transition records, and accepts precisely when the leftover suffix consists entirely of true bits, returning that suffix.
The two-argument skipping reader for one five-field transition record is primitive recursive jointly in the state bound \(n\) and the string. The record consists of four two-bit fields — the first, third, and fourth checked against a reserved block, the second unconstrained — followed by a halt-or-successor field whose unary index must be smaller than \(n + 1\).
The two-argument skipping reader that, given a bound \(n\) and a Boolean string, consumes a unary-encoded index (a block of true bits closed by a false bit) and returns the unconsumed suffix precisely when the index is smaller than \(n\), is primitive recursive jointly in \(n\) and the string.
For every fixed test \(v\) on pairs of Booleans, the skipping reader that consumes the first two bits \(a, b\) of a Boolean string and returns the remaining suffix when \(v(a, b)\) holds — rejecting otherwise, and rejecting strings with fewer than two bits — is primitive recursive.
The two-argument skipping reader for a halt-or-successor field is primitive recursive jointly in the bound and the string. This reader consumes one tag bit; on a true tag it further consumes a unary index checked to be smaller than the bound, and on a false tag it returns the rest of the string unchanged.
The unary reader on Boolean strings — the partial function that consumes a maximal leading block of true bits together with the terminating false bit and returns the pair consisting of the block’s length and the unconsumed suffix, undefined when the string contains no false bit — is primitive recursive.
Given a state-count parameter \(n\), a reader for one transition record of a one-work-tape machine over the binary alphabet with \(n + 1\) states. It reads, in order: an input-head movement, an optional work-tape write, a work-head movement, an optional output bit, and a successor-state field over \(n + 1\) states, failing as soon as any of the five fields fails, and assembles the results into a single transition record together with the unconsumed suffix.
For every transition record \(a\) of a one-work-tape binary machine with \(n + 1\) states and every bit string \(x\), the transition-record reader applied to the serialization of \(a\) (the dictionary codes of its input-head movement, work-tape write, work-head movement, and output, followed by its successor-state field) concatenated with \(x\) succeeds and returns \((a, x)\).
If the transition-record reader with state-count parameter \(n\) succeeds on a bit string \(x\), returning a record \(a\) and a suffix \(r\), then \(x\) is exactly the serialization of \(a\) — the codes of its five fields concatenated in order — followed by \(r\).
Given a bound \(n\), the reader that decodes a unary numeral (ones terminated by a zero) from a bit string and succeeds only when the decoded value \(i\) satisfies \(i \lt n\), in which case it returns \(i\) as an element of \(\{ 0, \dots , n-1\} \) together with the unconsumed suffix.
For a bound \(n\), an index \(i \lt n\), and any bit string \(x\): the range-checked unary reader with bound \(n\), applied to the self-delimiting unary code of \(i\) (\(i\) ones followed by a zero) concatenated with \(x\), succeeds and returns \((i, x)\).
If the range-checked unary reader with bound \(n\) succeeds on a bit string \(x\), returning an index \(i \lt n\) and a suffix \(r\), then \(x\) is exactly the self-delimiting unary code of \(i\) (\(i\) ones and a closing zero) followed by \(r\).
A reader decoding a fixed two-bit dictionary for an optional output bit: a leading \(00\) denotes producing no output, \(10\) outputting the bit \(0\), and \(11\) outputting the bit \(1\); a leading \(01\), or an input of fewer than two bits, is rejected. On success the decoded value is returned with the remaining suffix.
For every optional output bit \(b\) (no output, output \(0\), or output \(1\)) and every bit string \(x\), the optional-output reader applied to the two-bit dictionary code of \(b\) (\(00\), \(10\), or \(11\) respectively) followed by \(x\) succeeds and returns \((b, x)\).
A reader decoding a fixed two-bit dictionary for a head movement: a leading \(11\) denotes moving left, \(00\) staying put, and \(10\) moving right, in each case returning the movement together with the remaining suffix; a leading \(01\), or an input of fewer than two bits, is rejected.
For every head movement \(s\) (left, stay, or right) and every bit string \(x\), the head-movement reader applied to the two-bit dictionary code of \(s\) (\(11\), \(00\), or \(10\) respectively) followed by \(x\) succeeds and returns \((s, x)\).
Given a bound \(n\), a reader for a successor-state field: a leading \(0\) denotes halting (no successor state) and consumes only that bit, while a leading \(1\) is followed by a unary index that must be smaller than \(n\) and is returned as the live successor state, in each case together with the unconsumed suffix.
For a bound \(n\), a successor-state field \(s\) (either halt or a live state \(i \lt n\)), and any bit string \(x\): the successor-state reader with bound \(n\), applied to the encoding of \(s\) (a single \(0\) for halt; a \(1\) followed by the unary code of \(i\) for a live state) followed by \(x\), succeeds and returns \((s, x)\).
If the successor-state reader with bound \(n\) succeeds on a bit string \(x\), returning a field \(s\) and a suffix \(r\), then \(x\) is exactly the encoding of \(s\) (a single \(0\) for halt, or a \(1\) followed by the unary code of a live state below \(n\)) followed by \(r\).
Given a field reader producing values in a type \(A\), the reader that decodes three consecutive \(A\)-entries and returns the function assigning them to the three tape symbols blank, \(0\), \(1\) in that fixed order, together with the unconsumed suffix; it fails as soon as any of the three reads fails.
Let a field reader and a field writer for values in a type \(A\) be given such that reading from any written value followed by any suffix returns exactly that value and suffix. Then for every function \(f\) from the three tape symbols (blank, \(0\), \(1\)) to \(A\) and every bit string \(x\), the three-symbol row reader applied to the concatenation of the written values \(f(\mathrm{blank})\), \(f(0)\), \(f(1)\) in that order, followed by \(x\), succeeds and returns \((f, x)\).
Let a field reader and a field writer for values in a type \(A\) be such that every successful read on a string exhibits that string as the written value followed by the returned suffix. Then whenever the three-symbol row reader succeeds on a bit string \(x\), returning a row \(f\) and a suffix \(r\), the string \(x\) is exactly the concatenation of the written values \(f(\mathrm{blank})\), \(f(0)\), \(f(1)\) followed by \(r\).
Let \(M\) be a coded machine with state-count parameter \(n\). Applying the table reader — the length-\((n+1)\) vector reader whose entries are three-by-three grids of transition records indexed by the input and work symbols — to the concatenated serializations of all of \(M\)’s transition records (states in order, input and work symbols each in blank/\(0\)/\(1\) order) followed by any bit string \(x\) succeeds, returning the function that assigns to each state, input symbol, and work symbol the corresponding transition of \(M\), together with \(x\) untouched.
For every \(n \in \mathbb {N}\) and every bit string \(x\), the unary numeral reader applied to the string consisting of \(n\) ones, then a zero, then \(x\), succeeds and returns the pair \((n, x)\); in particular the suffix \(x\) is left untouched.
Given a field reader producing values in a type \(A\) and a length \(n\), the reader that decodes \(n\) consecutive \(A\)-entries and returns them as a vector indexed by \(\{ 0, \dots , n-1\} \), together with the unconsumed suffix; it fails as soon as any entry fails.
Let a field reader and a field writer for values in a type \(A\) be given such that reading from any written value followed by any suffix returns exactly that value and suffix. Then for every \(n\), every vector \(f\) indexed by \(\{ 0, \dots , n-1\} \), and every bit string \(x\), the length-\(n\) vector reader applied to the concatenation of the written entries \(f(0), \dots , f(n-1)\) in order, followed by \(x\), succeeds and returns \((f, x)\).
Let a field reader and a field writer for values in a type \(A\) be such that every successful read on a string exhibits that string as the written value followed by the returned suffix. Then whenever the length-\(n\) vector reader succeeds on a bit string \(x\), returning a vector \(f\) and a suffix \(r\), the string \(x\) is exactly the in-order concatenation of the written entries \(f(0), \dots , f(n-1)\) followed by \(r\).
A reader decoding a fixed two-bit dictionary for an optional work-tape write, whose possible values are: write nothing, write the blank symbol, write the bit \(0\), or write the bit \(1\), denoted by the leading pairs \(00\), \(01\), \(10\), \(11\) respectively. Inputs of fewer than two bits are rejected; on success the decoded value is returned with the remaining suffix.
For every optional work-tape write \(b\) (write nothing, write the blank symbol, write \(0\), or write \(1\)) and every bit string \(x\), the optional-write reader applied to the two-bit dictionary code of \(b\) (\(00\), \(01\), \(10\), or \(11\) respectively) followed by \(x\) succeeds and returns \((b, x)\).
The erased parser: it follows the same grammar as the machine parser — doubled-bit canonical state count \(n\), minimum-length guard \(81(n+1)\), a skipped unary initial state below \(n + 1\), then \((n+1) \cdot 3 \cdot 3\) skipped transition records, and an all-\(1\) remainder — but validates the fields without constructing any machine, returning only the unconsumed all-\(1\) suffix.
For every bit string \(x\), the result of the erased suffix scanner on \(x\) equals the result of the suffix-retaining machine parser on \(x\) with the machine component discarded: both fail together, and on success the scanner returns exactly the unconsumed suffix that the parser would return.
Given a state-count parameter \(n\), the reader that consumes one five-field transition record: four two-bit dictionary fields (input-head movement, work-tape write, work-head movement, output), each validated against its dictionary, followed by a successor-state field over \(n + 1\) states. It returns only the unconsumed suffix.
Let \(r\) be a partial reader on Boolean strings (returning either a suffix or failure), let \(n \in \mathbb {N}\), and let \(xs\) be a Boolean string. Applying \(r\) sequentially \(n\) times starting from \(xs\), threading each returned suffix into the next application and failing as soon as one application fails, gives the same result as the \(n\)-th iterate — evaluated at \(xs\) — of the total map on optional strings that feeds a present value to \(r\) and propagates failure.
A machine in code normal form consists of a natural number \(m\) (one less than the number of states, so the state space is never empty) together with a multi-tape Turing machine having exactly one work tape, binary tape alphabet \(\{ 0,1\} \), and state space the canonical finite type \(\{ 0, 1, \dots , m\} \).
The fixed, scheme-independent canonical serialization of a machine in code normal form: the pair encoding whose first component is the binary representation of the state-count parameter (self-delimiting via the doubled-bit region) and whose second component is the self-delimiting unary serialization of the initial state followed by the serialized transition records of the entire transition table, listed in a fixed enumeration order (states in increasing order, with the scanned input symbol and the scanned work symbol each ranging over blank, \(0\), \(1\)).
The finite-machine bundle associated with a machine in code normal form: the binary-alphabet machine package whose number of work tapes is \(1\), whose state type is the canonical finite state space \(\{ 0,\dots ,m\} \) of the coded machine, and whose transition table is that of the coded machine.
There exist a finite-state machine \(M\) over the binary alphabet and a constant \(c\) such that on every input string \(\alpha \) of length \(n\), the machine \(M\) halts within \(c \cdot (n + 1)\) steps with output the diagonal pair encoding \(\langle \alpha , \alpha \rangle \) (each bit of \(\alpha \) doubled, then the separator \(01\), then \(\alpha \) verbatim). In particular the diagonal pairing \(\alpha \mapsto \langle \alpha , \alpha \rangle \) is computable in linear time.
For every bit \(b\) and every bit string \(w\), the doubling of the string \(b\, w\) (where \(\mathrm{dbl}(v)\) writes each bit of \(v\) twice in succession) is \(b\, b\) followed by \(\mathrm{dbl}(w)\); that is, \(\mathrm{dbl}(b\, w) = b\, b\, \mathrm{dbl}(w)\).
An effective representation scheme: a representation scheme satisfying the algebraic laws, together with a canonizer — a finite-state binary-alphabet machine — and a time bound \(T : \mathbb {N} \to \mathbb {N}\) such that on every bit string \(\alpha \) the canonizer halts within \(T(|\alpha |)\) steps and outputs the fixed canonical serialization of the machine to which \(\alpha \) decodes.
Let \(S\) and \(H\) be sets of states, \(\iota : S \to H\), \(r \in H\), and \(p\) a Boolean word. For every action \(a\) and every configuration \(c\) of a Boolean Turing machine with \(k\) work tapes and state set \(S\) on an input \(x\), applying the forwarding transform of \(a\) (same head moves, writes and emission as \(a\), successor mapped through \(\iota \), a halt replaced by \(r\)) to the forwarding image of \(c\) (state mapped through \(\iota \), or \(r\) if \(c\) is halted, the same tapes and heads, output \(p\) followed by the output of \(c\)) gives the forwarding image of the configuration obtained by applying \(a\) to \(c\).
Let \(M\) and \(N\) be Turing machines over the Boolean alphabet with \(k\) work tapes each and state sets \(S\) and \(H\), and let \(\iota : S \to H\) and \(r \in H\). Suppose that for every \(s \in S\), every (possibly blank) input symbol \(\sigma \) and every tuple \(w\) of \(k\) work-tape symbols, the transition of \(N\) at \((\iota (s), \sigma , w)\) is the forwarding transform of the transition of \(M\) at \((s, \sigma , w)\): the same head moves, writes and emission, with the successor mapped through \(\iota \) and a halt replaced by entering \(r\). Let \(p\) be a Boolean word, \(c_0\) a configuration of \(M\) on an input \(x\), and \(t \in \mathbb {N}\) such that \(M\) run from \(c_0\) is not halted at any time \(t' \lt t\); then running \(N\) for \(t\) steps from the forwarding image of \(c_0\) yields the forwarding image of the configuration \(M\) reaches after \(t\) steps from \(c_0\), where the forwarding image of a configuration \(c\) has state \(\iota (q)\) (or \(r\) if \(c\) is halted), the input head, work tapes and work heads of \(c\), and output \(p\) followed by the output of \(c\).
Let \(S\) and \(H\) be sets of control states, \(\iota : S \to H\) a map and \(r \in H\) a designated return state. Given an action \(a\) of a Boolean Turing machine with \(k\) work tapes and state set \(S\), its forwarding transform is the action of a machine with the same \(k\) work tapes and state set \(H\) that has the same input-head move, the same writes and head moves on every work tape, and the same emitted bit (if any) as \(a\), and whose successor state is \(\iota (q')\) if \(a\) has successor state \(q'\), and \(r\) if \(a\) halts.
Let \(S\) and \(H\) be sets of control states, \(\iota : S \to H\) a map, \(r \in H\) a return state, and \(p\) a Boolean word. For a configuration \(c\) of a Boolean Turing machine with \(k\) work tapes and state set \(S\) on an input \(x\), its forwarding image is the configuration on \(x\) of a machine with \(k\) work tapes and state set \(H\) whose control state is \(\iota (q)\) if \(c\) is in state \(q\) and \(r\) if \(c\) is halted, whose input-head position, work-tape contents and work-head positions are those of \(c\), and whose output tape holds \(p\) followed by the output of \(c\).
Let \(z\), \(a\), \(b\) be bit strings, and consider the aligned pair parser, which reads aligned two-bit blocks from the left, collects the data bits \(0\) and \(1\) from the blocks \(00\) and \(11\), and at the first aligned block \(01\) returns the collected bits together with the remaining suffix. If the parser succeeds on \(z\) and returns \((a, b)\), then \(z\) is the pair encoding \(\langle a, b\rangle \), i.e. \(a\) with every bit written twice, then the separator bits \(0\) then \(1\), then \(b\) verbatim.
Let \(M\) be a machine over the binary alphabet with finitely many states and exactly one work tape. Then there exists a machine \(M'\) in code normal form such that for every input string \(x\), every output string \(w\), and every time bound \(t\), the bundled machine of \(M'\) halts on \(x\) within \(t\) steps with output \(w\) if and only if \(M\) does.
There exists an effective machine-code scheme: an encoding of one-work-tape binary machines in code normal form as Boolean strings, and a total decoding, such that a machine’s code followed by any number of trailing true padding bits decodes back to that machine; together with an actual finite binary Turing machine — the canonizer — and a time bound under which that machine computes, on every Boolean string, the fixed serialization of the machine the string decodes to.
For an alphabet \(\Sigma \), a finite nondeterministic Turing machine consists of a number \(k \in \mathbb {N}\) of work tapes, a finite state set with decidable equality, and a nondeterministic multi-tape Turing machine with \(k\) work tapes and those states, whose tape cells hold either a symbol of \(\Sigma \) or a blank.
Let \(N\) be a nondeterministic multi-tape Turing machine over the binary alphabet with finitely many states and two transition functions, one for each choice bit; let \(x\) be a binary string and \(t \in \mathbb {N}\). We say that \(N\) accepts \(x\) within \(t\) steps if there is a choice word \(w \in \{ 0,1\} ^t\) such that running \(N\) from its initial configuration on input \(x\) for \(t\) steps, using the \(i\)-th bit of \(w\) to select the transition function at step \(i\), reaches a halted configuration whose output is exactly the one-bit string \(1\). A branch halted with any other output, including the empty string and \(0\), is not accepting.
Let \(N\) be a finite nondeterministic multi-tape Turing machine over the binary alphabet, let \(x\) be a binary string, and let \(t \le t'\) be natural numbers. If some choice word of length \(t\) drives \(N\) on input \(x\) to a halted configuration with output exactly \(1\), then some choice word of length \(t'\) does as well; that is, if \(N\) accepts \(x\) within \(t\) steps then \(N\) accepts \(x\) within \(t'\) steps.
Let \(N\) be a finite nondeterministic multi-tape Turing machine over the binary alphabet, let \(L\) be a language of binary strings, and let \(T : \mathbb {N} \to \mathbb {N}\). We say that \(N\) decides \(L\) within time \(T\) if for every binary input \(x\): every choice word of length \(T(|x|)\) drives \(N\) from its initial configuration on \(x\) to a halted configuration, and \(x \in L\) if and only if some choice word of length \(T(|x|)\) drives \(N\) on \(x\) to a halted configuration with output exactly the one-bit string \(1\).
For an alphabet \(\Sigma \), a finite Turing machine consists of a number \(k \in \mathbb {N}\) of work tapes, a finite state set with decidable equality, and a deterministic multi-tape Turing machine with \(k\) work tapes and those states, whose tape cells hold either a symbol of \(\Sigma \) or a blank. This is the machine model of Arora–Barak §1.2, up to the declared model variations (append-only output tape, start-marker-free initialization).
[AB09, Claim 1.5]. Let \(\Gamma \) be a finite alphabet, \(e : \{ 0,1\} \hookrightarrow \Gamma \) an embedding of the two bits, \(M\) a Turing machine over \(\Gamma \) (with a read-only input tape, a write-only output tape, and \(k\) work tapes), \(f\) a function on binary strings, and \(T : \mathbb {N} \to \mathbb {N}\). Suppose \(M\) computes \(f\) via \(e\) within time \(T\): for every binary string \(x\), the run of \(M\) on the letterwise \(e\)-image of \(x\) halts within \(T(\left\lvert x\right\rvert )\) steps with output the letterwise \(e\)-image of \(f(x)\). Then there exist a constant \(c \in \mathbb {N}\) and a Turing machine \(M'\) over the binary alphabet with the same number \(k\) of work tapes such that \(M'\) computes \(f\) within time \(n \mapsto c\, (T(n)+1)\).
For any split \(x = p s\) of the input, there is a time \(t \le \left\lvert s\right\rvert + 1\) such that after \(t\) steps from the scan configuration at position \(\left\lvert p\right\rvert \) with empty output, the marker-existence scanner has halted with output the single bit recording whether \(s\) contains a \(1\).
For an embedding \(e : \{ 0,1\} \hookrightarrow \Gamma \) of the two bits into an alphabet \(\Gamma \) and a symbol \(a \in \Gamma \), the bit that is \(1\) exactly when \(a = e(1)\); in particular every symbol outside the image of \(e\) decodes to \(0\).
Let \(e : \{ 0,1\} \hookrightarrow \Gamma \) be an embedding of the bits into an alphabet \(\Gamma \). For every bit \(b\), applying the total bit decoder (which sends \(a \in \Gamma \) to \(1\) exactly when \(a = e(1)\)) to \(e(b)\) returns \(b\).
Given an injective block code \(E\) of \(\Gamma \cup \{ \text{blank}\} \) into blocks of \(N+1\) optional bits, a vector assigning to each of \(k\) work tapes a scanned element of \(\Gamma \cup \{ \text{blank}\} \), and a count \(j \in \mathbb {N}\): the per-tape buffer after \(j\) reading steps, namely for each tape \(i\) the block whose entries at offsets \(\lt j\) agree with the code of the scanned symbol of tape \(i\) and whose remaining entries are blank.
Let \(E\) be an injective block code of \(\Gamma \cup \{ \text{blank}\} \) into blocks of \(N+1\) optional bits, \(e : \{ 0,1\} \hookrightarrow \Gamma \) a bit embedding, \(x\) a binary string, and \(c\) a configuration of a \(\Gamma \)-machine with \(k\) work tapes on the letterwise \(e\)-image of \(x\). The encoded simulator configuration on the genuine binary input \(x\): its state is the initial reading state for the state of \(c\) (halted when \(c\) is halted); the input-head position is unchanged; work tape \(i\) is the binary tape induced by coding the corresponding tape of \(c\) blockwise through \(E\); work head \(i\) sits at offset \(0\) of the block of its source head position; and the output of \(c\) is decoded bitwise (a symbol \(a\) becoming \(1\) exactly when \(a = e(1)\)).
Let \(M\) be a Turing machine over a finite alphabet \(\Gamma \) and \(c\) a configuration of \(M\) on the letterwise \(e\)-image of a binary string. Running the binary block-code simulator for \(3(N+1)+2\) steps from the encoded configuration of \(c\) yields the encoded configuration of the result of one step of \(M\) from \(c\); when \(c\) is halted both encoded configurations are fixed points, so the identity holds in that case as well.
Suppose the injective block code \(E\) sends the blank to the all-blank block. Then for every Turing machine \(M\) over a finite alphabet \(\Gamma \), every bit embedding \(e : \{ 0,1\} \hookrightarrow \Gamma \), and every binary string \(x\), the initial configuration of the binary block-code simulator on input \(x\) equals the encoded initial configuration of \(M\) on the letterwise \(e\)-image of \(x\).
Let \(e : \{ 0,1\} \hookrightarrow \Gamma \) be a bit embedding, \(x\) a binary string, and \(c\) a configuration of a \(\Gamma \)-machine on the letterwise \(e\)-image of \(x\). The input symbol scanned by the encoded simulator configuration of \(c\) (which reads the genuine binary input \(x\) at the same input-head position), mapped through \(e\), equals the input symbol scanned by \(c\); in particular both are absent for the same head positions.
Let \(M\) be a Turing machine over a finite alphabet \(\Gamma \) and \(c\) a configuration of \(M\) on the letterwise \(e\)-image of a binary string. For every \(t \in \mathbb {N}\), running the binary block-code simulator for \((3(N+1)+2)\cdot t\) steps from the encoded configuration of \(c\) yields the encoded configuration of the configuration reached by \(M\) after \(t\) steps from \(c\).
For a finite alphabet \(\Gamma \), an injection of \(\Gamma \cup \{ \text{blank}\} \) into blocks of \(\left\lvert \Gamma \right\rvert +1\) optional bits: the blank maps to the all-blank block, and a symbol \(x \in \Gamma \) maps to the block whose entry at offset \(j\) is the genuine bit that is \(1\) exactly when \(j\) is the index of \(x\) under a fixed enumeration of \(\Gamma \).
Given an injection \(E\) of \(\Gamma \cup \{ \text{blank}\} \) into blocks of \(N+1\) optional bits, a total decoding map from blocks to \(\Gamma \cup \{ \text{blank}\} \) chosen as a left inverse of \(E\): every block in the image of \(E\) is sent to its unique preimage.
Let \(E\) be an injection of \(\Gamma \cup \{ \text{blank}\} \) into blocks of \(N+1\) optional bits, and let the decoder be the associated left inverse of \(E\). Then for every \(a \in \Gamma \cup \{ \text{blank}\} \), decoding the block \(E(a)\) returns \(a\).
Let \(M\) be a Turing machine over an alphabet \(\Gamma \) and \(e : \{ 0,1\} \hookrightarrow \Gamma \) a bit embedding, and suppose \(M\) computes the string function \(f : \{ 0,1\} ^* \to \{ 0,1\} ^*\) via \(e\) within time \(T\): for every binary string \(x\), the run of \(M\) on the letterwise \(e\)-image of \(x\) halts within \(T(\left\lvert x\right\rvert )\) steps with output the letterwise \(e\)-image of \(f(x)\). If, in the run of \(M\) on the \(e\)-image of some binary input \(x\), the transition taken from the configuration reached after \(t\) steps emits a symbol \(a \in \Gamma \), then \(a = e(b)\) for some bit \(b\).
The simulator configuration in the move sweep at counter \(r \in \{ 0,\dots ,N+1\} \), for a source configuration \(c\) and a source action \(a\): the work tapes, input position, and output are those of the encoded configuration of the result of applying \(a\) to \(c\); the state carries the pending record of \(a\) with move counter \(r\); and work head \(i\) sits at position \((N+1)p_i + (N+1-r)\, d_i\), where \(p_i\) is the source head position of \(c\) on tape \(i\) and \(d_i \in \{ -1,0,+1\} \) is the movement of \(a\) on tape \(i\).
One transition of the binary block-code simulator applied to the move-sweep configuration (source configuration \(c\), source action \(a\)) at counter \(0\) yields the encoded configuration of the result of applying \(a\) to \(c\); in particular the simulator halts exactly when the next state of \(a\) is absent.
For every counter \(r \le N+1\), running the binary block-code simulator for \(r+1\) steps from the move-sweep configuration (source configuration \(c\), source action \(a\)) at counter \(r\) reaches the encoded configuration of the result of applying \(a\) to \(c\).
For a counter \(r \neq 0\), one transition of the binary block-code simulator applied to the move-sweep configuration (source configuration \(c\), source action \(a\)) at counter \(r\) yields the move-sweep configuration at counter \(r-1\): each work head advances one cell in its pending direction.
For an alphabet \(\Gamma \), a state type \(S\), and \(k \in \mathbb {N}\) work tapes, a pending record consists of an optional next state (absence meaning halt), for each of the \(k\) work tapes an element of \(\Gamma \cup \{ \text{blank}\} \) to be written, and for each work tape a head movement direction in \(\{ -1,0,+1\} \).
Given the vector of symbols currently scanned under the \(k\) work heads and a source machine action \(a\) (next state, per-tape optional write and head move, input-head move, optional emission), the pending record retaining: the next state of \(a\); for each work tape the effective new symbol (the symbol \(a\) writes, or the scanned symbol when \(a\) writes nothing); and the head movement direction of \(a\) on each work tape.
The simulator configuration \(j\) steps into the read sweep of a macro step, for a source configuration \(c\), a source state \(q\), and \(j \in \{ 0,\dots ,N+1\} \): it agrees with the encoded configuration of \(c\) except that its state is the reading state with source state \(q\), counter \(j\), and the \(j\)-step buffers of the symbols scanned by \(c\), and each work head sits at offset \(j\) of the block of its source head position.
Let \(M\) be a Turing machine over a finite alphabet \(\Gamma \), \(c\) a configuration of \(M\) on the letterwise \(e\)-image of a binary string, and \(q\) a source state. One transition of the binary block-code simulator applied to the read-sweep configuration of \(c\) at the full counter \(N+1\) (all blocks buffered) yields the write-sweep configuration at counter \(N\) for the action produced by the transition function of \(M\) on the state \(q\), the input symbol scanned by \(c\), and the work-tape symbols scanned by \(c\): the buffers are decoded, the input head moves and any emission is bit-decoded and emitted, and each work head steps left onto the last entry of its block.
For a counter \(j \le N\), one transition of the binary block-code simulator applied to the read-sweep configuration of a source configuration \(c\) with source state \(q\) at counter \(j\) yields the read-sweep configuration at counter \(j+1\): the scanned code bits enter the buffers and every work head moves one cell right.
For a counter \(j \le N\), the work-tape symbols scanned by the read-sweep configuration of a source configuration \(c\) at counter \(j\) are, on each tape \(i\), the entry at offset \(j\) of the code of the symbol scanned by \(c\) on tape \(i\).
The finite state space of the block-code simulator with parameters \(\Gamma \), \(S\), \(k\), \(N\): either a reading state, consisting of a source state \(q \in S\), a counter \(j \in \{ 0,\dots ,N+1\} \), and for each of the \(k\) work tapes a partially filled buffer block of \(N+1\) optional bits; or a pending record (optional next state, per-tape symbol to write, per-tape direction) together with either a write counter in \(\{ 0,\dots ,N\} \) or a move counter in \(\{ 0,\dots ,N+1\} \).
Given an injective block code \(E\) of \(\Gamma \cup \{ \text{blank}\} \) into blocks of \(N+1\) optional bits, a source tape \(t : \mathbb {Z} \to \Gamma \cup \{ \text{blank}\} \), a position \(p \in \mathbb {Z}\), a symbol \(w\) to be written at \(p\), and a counter \(r \in \mathbb {N}\): the family of blocks coding the partially rewritten tape, where every block \(z \neq p\) is the code of \(t(z)\), and block \(p\) takes its entries at offsets \(\ge r\) from the code of \(w\) and its entries at offsets \(\lt r\) from the code of \(t(p)\).
For \(j \in \{ 0,\dots ,N\} \), overwriting the binary tape induced by the counter-\((j+1)\) partially rewritten block family (tape \(t\), write position \(p\), new symbol \(w\)) at tape position \((N+1)p + j\) with entry \(j\) of the code of \(w\) yields the binary tape induced by the counter-\(j\) family: blocks other than \(p\) are unchanged, and within block \(p\) the only new entry is the one at offset \(j\).
At counter \(0\) (after the write is complete) the partially rewritten block family for tape \(t\), write position \(p\), and new symbol \(w\) is the blockwise coding of the updated tape that agrees with \(t\) everywhere except that position \(p\) now holds \(w\).
Given \(N \in \mathbb {N}\) and a family \(F\) assigning to every integer \(z\) a block of \(N+1\) optional bits, the induced binary tape \(\mathbb {Z} \to \{ 0,1\} \cup \{ \text{blank}\} \): position \(p\) holds entry \(p \bmod (N+1)\) of the block \(F(\lfloor p/(N+1) \rfloor )\).
Let \(F\) assign to every integer a block of \(N+1\) optional bits. The binary tape induced by \(F\) (position \(p\) holding entry \(p \bmod (N+1)\) of block \(F(\lfloor p/(N+1)\rfloor )\)) reads, at position \((N+1)z + j\) with \(z \in \mathbb {Z}\) and \(j \in \{ 0,\dots ,N\} \), exactly the entry \(j\) of \(F(z)\).
Let \(F\) assign to every integer a block of \(N+1\) optional bits, let \(z \in \mathbb {Z}\), \(j \in \{ 0,\dots ,N\} \), and let \(b\) be an optional bit. Overwriting the binary tape induced by \(F\) at position \((N+1)z + j\) with the value \(b\) yields the binary tape induced by the family that agrees with \(F\) everywhere except that entry \(j\) of block \(z\) becomes \(b\).
Given a finite alphabet \(\Gamma \), an injective block code \(E\) of \(\Gamma \cup \{ \text{blank}\} \) into blocks of \(N+1\) optional bits, a bit embedding \(e : \{ 0,1\} \hookrightarrow \Gamma \), and a Turing machine \(M\) over \(\Gamma \) (read-only input tape, write-only output tape, \(k\) work tapes), the simulating machine over the binary alphabet with the same number \(k\) of work tapes. Its control cycles through: \(N+1\) reading steps that copy the block under each work head into a buffer while moving right; one dispatch step that decodes the buffers, applies the transition of \(M\) (with the scanned input bit mapped through \(e\)), performs the input-head move of \(M\), emits the bit-decoded image of any emitted symbol, stores the resulting pending record, and moves each work head one cell left; \(N+1\) writing steps that write the coded new symbols right to left; and \(N+2\) moving steps that shift each work head one whole block in its source direction and then re-enter the reading state (halting if the source transition halts).
Let \(\Gamma \) be a finite alphabet, \(E\) an injective block code of \(\Gamma \cup \{ \text{blank}\} \) into blocks of \(N+1\) optional bits that sends the blank to the all-blank block, \(e : \{ 0,1\} \hookrightarrow \Gamma \) a bit embedding, and \(M\) a Turing machine over \(\Gamma \). If, started on the letterwise \(e\)-image of a binary string \(x\), \(M\) halts within \(t\) steps with output the letterwise \(e\)-image of a binary string \(w\), then the binary block-code simulator, started on the input \(x\) itself, halts within \((3(N+1)+2)\cdot t\) steps with output \(w\).
The simulator configuration in the write sweep at counter \(j \in \{ 0,\dots ,N\} \), for a source configuration \(c\) and a source action \(a\): the state carries the pending record of \(a\) with write counter \(j\); work tape \(i\) is the binary tape induced by the partially rewritten block family in which, at the block of the source head position, the entries at offsets \(\gt j\) already hold the code of the effective new symbol (the symbol \(a\) writes, or the currently scanned symbol when \(a\) writes nothing); work head \(i\) sits at offset \(j\) of that block; and the input position and output are those of the encoded configuration of the result of applying \(a\) to \(c\).
One transition of the binary block-code simulator applied to the write-sweep configuration (source configuration \(c\), source action \(a\)) at counter \(0\) writes the last code entry in place and yields the move-sweep configuration at counter \(N+1\): the work tapes now hold the complete blockwise coding of the tapes of the result of applying \(a\) to \(c\).
For a counter \(j \neq 0\), one transition of the binary block-code simulator applied to the write-sweep configuration (source configuration \(c\), source action \(a\)) at counter \(j\) yields the write-sweep configuration at counter \(j-1\): on each tape the entry at offset \(j\) of the code of the effective new symbol is written and the head moves one cell left.
The blank block of \(k\) sweep cells: the cell for tape \(i\) carries index \(i\), no payload, a head flag equal to a given mark, and a cleared left-neighbour flag. With the mark set it is the initialized origin block; with the mark cleared it is a fresh guard block.
Given two Turing machines \(M_1\) and \(M_2\) over the Boolean alphabet with finitely many states, and a Boolean \(b\), the machine whose work tapes are the disjoint union of the tapes of \(M_1\) and \(M_2\), whose state type is the disjoint union of their state types, whose transitions run \(M_1\) on the left tape block from left states and \(M_2\) on the right tape block from right states, and whose initial state is the initial state of \(M_1\) when \(b\) is true and of \(M_2\) when \(b\) is false. Only the initial state depends on \(b\); the transition table does not.
Let \(M_1\) and \(M_2\) be Turing machines over the Boolean alphabet with finitely many states, let \(b\) be a Boolean, let \(x\) and \(w\) be Boolean words, and let \(t\) be a natural number. The branch union of \(M_1\) and \(M_2\) selected by \(b\) halts on input \(x\) within \(t\) steps with output \(w\) if and only if the selected machine — \(M_1\) when \(b\) is true, \(M_2\) when \(b\) is false — halts on input \(x\) within \(t\) steps with output \(w\).
Let \(M_1\) and \(M_2\) be machines over the Boolean alphabet with finitely many states, and suppose that \(M_1\), on input \(x\), halts within \(t_1\) steps with output \(y\). Then there exist a step count \(a \le t_1 + \left\lvert y\right\rvert + 2\), a position \(p\) of the physical input head over \(x\), and contents and head positions for the left tape block, such that running the buffered sequential composition of \(M_1\) and \(M_2\) for \(a\) steps from its initial configuration on \(x\) yields the second-phase embedding of the initial configuration of \(M_2\) on virtual input \(y\), with arrival tag true, physical position \(p\), and that inactive left-block data.
Given machines \(M_1\) and \(M_2\) over the Boolean alphabet with finitely many states, having \(k_1\) and \(k_2\) work tapes respectively, the machine with \(k_1 + (1 + k_2)\) work tapes that operates in phases. In the first phase it simulates \(M_1\) on the left tape block, diverting every symbol that \(M_1\) would output onto the middle buffer tape instead of the real output; the simulated halt of \(M_1\) leads to a live administrative rewind state, in which the buffer head scans left until it reaches the blank preceding the stored word. It then moves right and dispatches into the second phase, a simulation of \(M_2\) on the right tape block in which the input reads of \(M_2\) are served from the buffer, with movements clamped at the boundary blanks of the stored word by a Boolean arrival tag carried in the state. Only the second phase emits symbols to the real output or halts.
Let \(M_1\) and \(M_2\) be multi-tape Turing machines over the Boolean alphabet, let \(x\), \(y\), \(o\) be Boolean words, and let \(t_1, t_2 \in \mathbb {N}\). Consider the buffered sequential composition machine of \(M_1\) and \(M_2\), which simulates \(M_1\) on its input with \(M_1\)’s output diverted to a buffer work tape, rewinds the buffer, and then simulates \(M_2\) reading the buffer as its input. If \(M_1\) halts on input \(x\) within \(t_1\) steps with output \(y\), and \(M_2\) halts on input \(y\) within \(t_2\) steps with output \(o\), then this composition machine halts on input \(x\) within \(t_1 + \left\lvert y\right\rvert + 2 + t_2\) steps with output \(o\). Neither machine is required to halt on any other input.
Given machines \(M_1\) and \(M_2\) over the Boolean alphabet with finitely many states and a configuration \(c\) of \(M_1\) on input \(x\), the configuration of the buffered sequential composition of \(M_1\) and \(M_2\) on input \(x\) whose state is the first-phase image of the state of \(c\) (a live administrative state even when \(c\) has halted), whose input head agrees with \(c\), whose left tape block carries the work tapes and heads of \(c\), whose middle buffer tape stores the output of \(c\) contiguously with the buffer head on the blank immediately to its right, whose right tape block is blank with heads at the origin, and whose real output is empty.
For all machines \(M_1\) and \(M_2\) over the Boolean alphabet with finitely many states and every input \(x\), the initial configuration of the buffered sequential composition of \(M_1\) and \(M_2\) on \(x\) equals the first-phase embedding of the initial configuration of \(M_1\) on \(x\), in which the buffer stores the empty word.
Let \(M_1\) and \(M_2\) be machines over the Boolean alphabet with finitely many states, and let \(c\) be a halted configuration of \(M_1\) on input \(x\) with output \(w\). Running the buffered sequential composition of \(M_1\) and \(M_2\) for exactly \(\left\lvert w\right\rvert + 2\) steps from the first-phase embedding of \(c\) yields the second-phase embedding of the initial configuration of \(M_2\) on virtual input \(w\), with arrival tag true, the physical input head parked where \(c\) left it, and the work tapes and head positions of \(c\) retained as inactive left-block data. The real output remains empty throughout.
Let \(M_1\) and \(M_2\) be machines over the Boolean alphabet with finitely many states, let \(c\) be a configuration of \(M_1\) on input \(x\), and let \(t\) be a number of steps such that the run of \(M_1\) from \(c\) has not halted at any time strictly before \(t\). Then running the buffered sequential composition of \(M_1\) and \(M_2\) for \(t\) steps from the first-phase embedding of \(c\) yields the first-phase embedding of the run of \(M_1\) for \(t\) steps from \(c\). The time \(t\) itself may be the first halting time of \(M_1\).
Let \(M_1\) and \(M_2\) be machines over the Boolean alphabet with finitely many states, and let \(c\) be a configuration of \(M_1\) on input \(x\) that has not halted. One step of the buffered sequential composition of \(M_1\) and \(M_2\) from the first-phase embedding of \(c\) equals the first-phase embedding of one step of \(M_1\) from \(c\). In particular a symbol emitted by \(M_1\) on that step — including on its halting transition — lands on the buffer with the buffer head advancing past it, the real output stays empty, and a simulated halt of \(M_1\) leaves the composite in a live administrative state.
Given machines \(M_1\) and \(M_2\) over the Boolean alphabet with finitely many states, a Boolean word \(y\) stored on the buffer, a position \(p \in \{ 0,\dots ,\left\lvert x\right\rvert +1\} \) of the physical input head over the real input \(x\), arbitrary contents and head positions for the left tape block, and a natural number \(j\), the configuration of the buffered sequential composition of \(M_1\) and \(M_2\) on input \(x\) in the live administrative scan state, with the physical input head parked at \(p\), the left block carrying the given inactive data, the buffer storing \(y\) with its head at cell \(j - 1\), the right block blank with heads at the origin, and empty real output. The scan state is live even when \(j = 0\).
Let \(M_1\) and \(M_2\) be machines over the Boolean alphabet with finitely many states, let \(y\) be a Boolean word stored on the buffer, let \(p\) be a position of the physical input head over the real input \(x\), and fix arbitrary inactive contents and head positions for the left tape block. For every \(j \le \left\lvert y\right\rvert \), running the buffered sequential composition of \(M_1\) and \(M_2\) for exactly \(j + 1\) steps from the rewind-scan configuration with buffer head at cell \(j - 1\) yields the second-phase embedding of the initial configuration of \(M_2\) on virtual input \(y\), with arrival tag true, physical position \(p\), and the same inactive data. This includes the case of an empty word \(y\).
Given machines \(M_1\) and \(M_2\) over the Boolean alphabet with finitely many states, a configuration \(c\) of \(M_2\) on a virtual input \(y\), a Boolean tag \(b\), a position \(p \in \{ 0,\dots ,\left\lvert x\right\rvert +1\} \) of the physical input head over the real input \(x\), and arbitrary contents and head positions for the left tape block of \(M_1\), the configuration of the buffered sequential composition of \(M_1\) and \(M_2\) on input \(x\) whose state is the second-phase image of the state of \(c\) paired with \(b\) (halted exactly when \(c\) is halted), whose physical input head is parked at \(p\), whose left block carries the given inactive data, whose middle buffer stores \(y\) with the buffer head at the virtual input-head position of \(c\) minus one, whose right block carries the work tapes and heads of \(c\), and whose real output is the output of \(c\).
Let \(M_1\) and \(M_2\) be machines over the Boolean alphabet with finitely many states, let \(c\) be a configuration of \(M_2\) on a virtual input \(y\), and let \(b\) be a Boolean tag valid for the virtual input-head position of \(c\). Fix a position \(p\) of the physical input head over the real input \(x\), arbitrary inactive contents and head positions for the left tape block, and a number of steps \(t\). Then there exists a tag \(b'\), valid for the input-head position of the run of \(M_2\) for \(t\) steps from \(c\), such that running the buffered sequential composition of \(M_1\) and \(M_2\) for \(t\) steps from the second-phase embedding of \(c\) with tag \(b\) equals the second-phase embedding of that run with tag \(b'\), keeping \(p\) and the inactive data unchanged. Completed outputs and absorbing halting are thereby preserved.
Let \(M_1\) and \(M_2\) be machines over the Boolean alphabet with finitely many states, let \(c\) be a configuration of \(M_2\) on a virtual input \(y\), and let \(b\) be a Boolean tag valid for the virtual input-head position of \(c\) (forced false at the left boundary and true at the right boundary). Fix a position \(p\) of the physical input head over the real input \(x\) and arbitrary inactive contents and head positions for the left tape block. Then there exists a tag \(b'\), valid for the input-head position reached after one step of \(M_2\) from \(c\), such that one step of the buffered sequential composition of \(M_1\) and \(M_2\) from the second-phase embedding of \(c\) with tag \(b\) equals the second-phase embedding of one step of \(M_2\) from \(c\) with tag \(b'\), keeping \(p\) and the inactive data unchanged. The halted case is included: a halted \(c\) stays halted on both sides.
For a Boolean word \(w\), the work-tape content over the integer cells that stores \(w\) contiguously from cell zero: cell \(z\) holds the \(z\)-th symbol of \(w\) (zero-indexed) when \(0 \le z \lt \left\lvert w\right\rvert \), and is blank at every other integer cell.
Let \(w\) be a Boolean word and \(b\) a Boolean symbol. The buffer tape storing \(w\) followed by the single symbol \(b\) equals the buffer tape storing \(w\) updated at cell \(\left\lvert w\right\rvert \) to hold \(b\): exactly the cell that was the old right blank changes, and every other integer cell is unaffected.
Let \(c\) be a configuration of a machine with \(k\) work tapes over the Boolean alphabet whose input is the word \(w\), and let \(p \in \{ 0,\dots ,\left\lvert w\right\rvert +1\} \) be its input-head position. Cell \(p - 1\) of the buffer tape storing \(w\) holds exactly the symbol scanned by the input head of \(c\): the corresponding input symbol at interior positions, and a blank at the two boundary positions.
Let \(w\) be a Boolean word and \(i\) a natural number. Cell \(i\) of the buffer tape storing \(w\) holds the optional \(i\)-th entry of \(w\): the symbol \(w_i\) when \(i \lt \left\lvert w\right\rvert \) and a blank otherwise. In particular cell \(\left\lvert w\right\rvert \) is the blank immediately to the right of the stored word.
Let \(q\) be a control state on which the machine stalls (every configuration in state \(q\) is a fixed point of the step function), and suppose the run from \(c\) reaches a configuration \(d\) in state \(q\) after \(T\) steps. Then there is a time \(t \le T\) such that no configuration strictly before time \(t\) is in state \(q\) and the configuration at time \(t\) is already \(d\).
Every word \(v\) satisfies exactly one of: \(v\) contains no \(1\) and the last-marker split of \(v\) is undefined; or \(v\) contains a \(1\) and the last-marker split of \(v\) is some word \(u\) (the prefix of \(v\) strictly before its last \(1\)) such that moreover, for every prefix \(p\), the last-marker split of \(p v\) is \(p u\).
Let a machine compute the total function \(g\) within a monotone time bound \(T_g\). Then the buffered composition of the second-component pair extractor with that machine computes, within \(6(n+1) + T_g(n) + 1\) steps on inputs of length \(n\), the function sending \(x\) to \(g(b)\), where \(b\) is the second component of \(x\) under the self-delimiting pair code (the empty word when \(x\) is malformed).
Let \(q \gt 0\) be the unary side length, let \(i \le c\) be a loop level, and suppose every head at a level \(\le i\) is at displacement \(0\) except that head \(i\) sits at \(j\), where \(j + r = q\). Then exactly \(r\, (T_i + 2) + q + 2\) steps (with \(T_i\) the exact level-\(i\) nest cost) execute the remaining \(r\) iterations of loop \(i\), append \(r \cdot C q^i\) ones to the output, reset head \(i\) to \(0\), and enter the advance state of level \(i+1\); heads above level \(i\) are arbitrary and remain unchanged.
From the rewind state of loop level \(i\) with head \(i\) at displacement \(j-1\), for any \(j \le q\), exactly \(j+1\) steps reset head \(i\) to displacement \(0\) and enter the advance state of level \(i+1\); the other loop heads and the accumulated output are unchanged.
For every input \(x\), running the generator for exactly \(2(\left\lvert x\right\rvert +1)\) steps from its initial configuration yields the configuration in which every work tape holds the unary word of length \(\left\lvert x\right\rvert +1\), the control is in the outermost loop state, all heads are at displacement \(0\), and the output is empty; the extra unary cell makes empty input need no special case.
For all parameters \(c\) and \(C\), the unary polynomial generator computes the function sending a word of length \(n\) to the run of \(C\, (n+1)^{c+1}\) ones, within \((C + 5(c+1) + 4)\, (n+1)^{c+1}\) steps; this covers coefficient \(C = 0\) and empty input.
For parameters \(c\) and \(C\), the control-state type of the unary polynomial generator has a copy state, a setup state, a loop state and a rewind state for each of the \(c+1\) loop levels, an advance state for each of \(c+2\) indices (one per level and one past the last), and an emission state for each remaining count from \(0\) to \(C\).
Given a work-tape index \(i\), a direction \(d\), and a control state \(s\), the head-only move action is the action that moves work head \(i\) one cell in direction \(d\), keeps the input head and all other work heads in place, writes no tape cell, emits no output, and enters \(s\).
For every work-tape index \(i\) and direction \(d\), applying the head-only move action for \(i\), \(d\), and a target state to a generator loop configuration yields the loop configuration in that target state whose \(i\)-th work-head position is increased by \(d\); all tapes, the input position, the other heads, and the output are unchanged.
A machine with \(c+1\) work tapes that writes a run of \(C\, (n+1)^{c+1}\) ones on its output tape, where \(n\) is the input length: in one input scan it copies \(n+1\) in unary onto every work tape in parallel, and then runs \(c+1\) nested loops, each of side \(n+1\), emitting \(C\) ones at each visit of the innermost body; a completed inner loop is rewound at a cost charged to the iterations that just completed.
For a fixed word \(w \in \{ 0,1\} ^*\), a finite Turing machine with no work tapes and \(\left\lvert w\right\rvert +1\) control states that first writes the bits of \(w\) to its output tape, one per step, and then copies its input to the output verbatim.
For every fixed word \(w\) and every input \(x\), the prefix emitter for \(w\) halts on \(x\) within \(\left\lvert w\right\rvert + \left\lvert x\right\rvert + 1\) steps, and its output is the concatenation \(w x\); the final step is the halting transition on the right blank.
For every \(i \le \left\lvert x\right\rvert \), after \(i\) steps from the configuration in which the whole fixed word \(w\) has been emitted and the input head sits at the start of \(x\), the prefix emitter has advanced the input head by \(i\) cells and its output is \(w\) followed by the first \(i\) bits of \(x\).
For every \(i \le \left\lvert w\right\rvert \), after \(i\) steps from its initial configuration on input \(x\) the fixed-prefix emitter for \(w\) is in control state \(i\), the input head has not moved, and the output is the first \(i\) bits of \(w\).
Suppose a machine has a start state that unconditionally moves the input head left into a scanning state, and the scanning state keeps moving left on symbols and exits right into a destination state on the left blank. Then from any configuration \(c\) in the start state there is a time \(r \le i + 2\), where \(i\) is the current input position, after which the machine is in the destination state with the input head back at the first cell, all tapes and the output unchanged.
A finite Turing machine \(M\) over an alphabet \(\Sigma \) computes the string function \(f \colon \Sigma ^{*} \to \Sigma ^{*}\) if for every input word \(x\) there exists a number of steps \(t\) such that \(M\) halts on \(x\) within \(t\) steps with \(f(x)\) written on its output tape. This is the notion of computability underlying the uncomputability results (Arora–Barak §1.4–1.5).
Let \(\Sigma \) be a finite alphabet, let \(M\) be a finite Turing machine over \(\Sigma \), and let \(f\) be a total function from strings over \(\Sigma \) to strings over \(\Sigma \). If \(M\) computes \(f\) — on every input \(x\) it halts, after some finite number of steps, with \(f(x)\) on its output tape, with no stated time bound — then there exists a function \(T : \mathbb {N} \to \mathbb {N}\) such that \(M\) computes \(f\) within time \(T\), i.e. on every input \(x\) the machine \(M\) halts with output \(f(x)\) within \(T(\left\lvert x\right\rvert )\) steps, where \(\left\lvert x\right\rvert \) is the length of \(x\). No monotonicity, positivity, or constructibility of \(T\) is claimed.
Let \(M\) be a finite Turing machine over an alphabet \(\Sigma \) that computes the string function \(g \colon \Sigma ^{*} \to \Sigma ^{*}\), and let \(x\) and \(w\) be words over \(\Sigma \). Then there exists a number of steps \(t\) such that \(M\), started on \(x\), halts within \(t\) steps with \(w\) on its output tape if and only if \(w = g(x)\).
A finite Turing machine \(M\) over an alphabet \(\Sigma \) computes the string function \(f \colon \Sigma ^{*} \to \Sigma ^{*}\) in time \(T \colon \mathbb {N} \to \mathbb {N}\) if for every input word \(x\), the machine \(M\) halts on \(x\) within \(T(\left\lvert x\right\rvert )\) steps with \(f(x)\) written on its output tape (Arora–Barak, Definition 1.3).
Let \(M\) be a finite Turing machine over an alphabet \(\Sigma \). If \(M\) computes the string function \(f \colon \Sigma ^{*} \to \Sigma ^{*}\) in time \(T \colon \mathbb {N} \to \mathbb {N}\) (halting within \(T(\left\lvert x\right\rvert )\) steps on every input \(x\) with \(f(x)\) on its output tape), then \(M\) computes \(f\), i.e. on every input it eventually halts with the value of \(f\) on its output tape.
For every fixed bit \(b\) there exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes the function \(x \mapsto x b\) (append the bit \(b\) to the input word) within \(c\, (n+1)\) steps on inputs of length \(n\).
Let \(M_1\) and \(M_2\) be multi-tape Turing machines over the Boolean alphabet, let \(f\) and \(g\) be functions on Boolean words, and let \(T_1, T_2 : \mathbb {N} \to \mathbb {N}\) with \(T_2\) monotone. Suppose \(M_1\) computes \(f\) in time \(T_1\) and \(M_2\) computes \(g\) in time \(T_2\), meaning that on every input \(x\) the machine halts within \(T_i(\left\lvert x\right\rvert )\) steps with the function value on its output tape. Then there exist a machine \(M\) and a constant \(c\) such that \(M\) computes the composite \(g \circ f\) in time \(n \mapsto c \cdot \bigl(T_1(n) + T_2(T_1(n)) + 1\bigr)\).
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet, let \(p\) be a Boolean predicate on Boolean words, \(f_1\) and \(f_2\) functions from Boolean words to Boolean words, and \(T_0, T_1, T_2 : \mathbb {N} \to \mathbb {N}\). Suppose that on every input \(x\), \(D\) halts within \(T_0(|x|)\) steps with output the one-bit word \(p(x)\), \(M_1\) halts within \(T_1(|x|)\) steps with output \(f_1(x)\), and \(M_2\) halts within \(T_2(|x|)\) steps with output \(f_2(x)\). Then there exist a finite-state Turing machine \(M\) over the Boolean alphabet and a constant \(c \in \mathbb {N}\) such that on every input \(x\), \(M\) halts within \(c \cdot (T_0(|x|) + \max (T_1(|x|), T_2(|x|)) + 1)\) steps with output \(f_1(x)\) if \(p(x)\) holds and \(f_2(x)\) otherwise.
For every fixed Boolean word \(w\) there exist a multi-tape Turing machine \(M\) over the Boolean alphabet and a constant \(c\) such that \(M\) computes the constant function \(x \mapsto w\) in time \(c \cdot (n + 1)\): on every input \(x\), within \(c \cdot (\left\lvert x\right\rvert + 1)\) steps \(M\) halts with output exactly \(w\).
There exist a multi-tape Turing machine \(M\) over the Boolean alphabet and a constant \(c\) such that \(M\) computes the identity function on Boolean words in time \(c \cdot (n + 1)\): on every input \(x\), within \(c \cdot (\left\lvert x\right\rvert + 1)\) steps \(M\) halts with output exactly \(x\).
For all fixed Boolean words \(w_0\), \(u\), \(v\) there exist a multi-tape Turing machine \(M\) over the Boolean alphabet and a constant \(c\) such that \(M\) computes in time \(c \cdot (n + 1)\) the function sending a Boolean word \(w\) to \(u\) when \(w = w_0\) and to \(v\) otherwise.
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)\) steps on inputs of length \(n\), the fixed-width binary successor of its input read least-significant bit first, with empty output on overflow (an all-ones input).
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes the function sending each input \(x\) to the little-endian binary representation of \(\left\lvert x\right\rvert \), within \(c\, (n+1)\) steps on inputs of length \(n\).
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)\) steps on inputs of length \(n\), the function sending a word that encodes the pair \((a,b)\) under the self-delimiting pair code to the concatenation \(a b\), and any malformed word to the empty word.
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)\) steps on inputs of length \(n\), the function sending \(x\) to the self-delimiting pair encoding of \((x,x)\): each bit of \(x\) doubled, the separator bits \(0,1\), then \(x\).
For every fixed word \(\alpha \) there exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes the function sending \(x\) to the self-delimiting pair encoding of \((\alpha , x)\) — each bit of \(\alpha \) doubled, the separator bits \(0,1\), then \(x\) — within \(c\, (n+1)\) steps on inputs of length \(n\).
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)\) steps on inputs of length \(n\), the function sending a word in the self-delimiting pair code to the first component of the encoded pair, and any malformed word to the empty word.
For all constants \(C\) and \(e\) there exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)^{e+1}\) steps on inputs of length \(n\), the one-bit function that on a word encoding the pair \((a,b)\) under the self-delimiting pair code outputs whether \(\left\lvert b\right\rvert \le C\, (\left\lvert a\right\rvert +1)^{e}\), and on any malformed word outputs \(0\).
Let a machine compute the total function \(g\) within a monotone time bound \(T_g\). Then there exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n + 1 + T_g(n))\) steps on inputs of length \(n\), the function sending a word that encodes the pair \((a,b)\) under the self-delimiting pair code to the pair encoding of \((a, g(b))\), and any malformed word to the empty word.
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)\) steps on inputs of length \(n\), the function sending a word in the self-delimiting pair code to the second component of the encoded pair, and any malformed word to the empty word.
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)\) steps on inputs of length \(n\), the one-bit function recording whether the input is a valid word of the self-delimiting pair code.
For all constants \(C\) and \(e\) there exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes the function sending each input of length \(n\) to the little-endian binary representation of \(C\, (n+1)^{e}\), within \(c\, (n+1)^{e+1}\) steps.
For all constants \(C\) and \(e\) there exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes the function sending each input of length \(n\) to the run of \(C\, (n+1)^{e}\) ones, within \(c\, (n+1)^{e+1}\) steps.
For every fixed word \(w \in \{ 0,1\} ^*\) there exist a finite Turing machine \(M\) and a constant \(c\) (which may depend on \(w\)) such that \(M\) computes the function \(x \mapsto w x\) (prepend \(w\) to the input) within \(c\, (n+1)\) steps on inputs of length \(n\).
For all constants \(C\) and \(e\) there exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)^{e+2}\) steps on inputs of length \(n\), the function that on input \(w\) searches for the least \(i \le \left\lvert w\right\rvert \) with \(i + C\, (i+1)^{e} = \left\lvert w\right\rvert \) and, on success, outputs the self-delimiting pair encoding of the first \(i\) bits of \(w\) paired with the rest of \(w\), and on failure outputs the empty word.
Let \(f : \mathbb {N} \to \mathbb {N}\) and let a machine \(E\) compute the function sending an input of length \(m\) to the little-endian binary representation of \(f(m)\) within a monotone time bound \(T_E\). Then there exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)(T_E(n+1) + n + 2)\) steps on inputs of length \(n\), the function that on input \(w\) finds the least \(i \le \left\lvert w\right\rvert \) with \(i + f(i) = \left\lvert w\right\rvert \) and, on success, outputs the self-delimiting pair encoding of the first \(i\) bits of \(w\) paired with the rest of \(w\), and on failure outputs the empty word.
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)^{2}\) steps on inputs of length \(n\), the following function: a word encoding the pair \((a,v)\) under the self-delimiting pair code, where \(v\) contains a \(1\) and \(u\) is the prefix of \(v\) strictly before its last \(1\), is sent to the pair encoding of \((a,u)\); a word whose second component contains no \(1\), or any malformed word, is sent to the empty word.
There exist a finite Turing machine \(M\) and a constant \(c\) such that \(M\) computes, within \(c\, (n+1)\) steps on inputs of length \(n\), the function splitting its input at the leading unary token — the maximal leading run of ones with its terminating zero, if present — and emitting the self-delimiting pair encoding of the token and the remainder.
Let \(\alpha \) and \(\Gamma \) be alphabets, let \(e \colon \alpha \hookrightarrow \Gamma \) be an injection, and let \(M\) be a finite Turing machine over \(\Gamma \). The machine \(M\) computes the string function \(f \colon \alpha ^{*} \to \alpha ^{*}\) via \(e\) in time \(T \colon \mathbb {N} \to \mathbb {N}\) if for every word \(x\) over \(\alpha \), on the input obtained by applying \(e\) letterwise to \(x\), the machine \(M\) halts within \(T(\left\lvert x\right\rvert )\) steps with the letterwise image of \(f(x)\) under \(e\) written on its output tape. This is how a machine over a larger alphabet is said to compute a function on a smaller one (Arora–Barak §1.3.1).
A finite Turing machine \(M\) over an alphabet \(\Sigma \) computes the output \(w\) from the input \(x\) within \(t\) steps if there exists a space bound \(s\) such that \(M\), started on \(x\), halts within \(t\) steps using space \(s\) with exactly \(w\) written on its output tape; that is, the space used is existentially quantified away from the time-and-space notion of computation (Arora–Barak, Definition 1.3).
Let \(M\) be a finite Turing machine over an alphabet \(\Sigma \), let \(x\) and \(w\) be words over \(\Sigma \), and let \(t \le t'\) be natural numbers. If \(M\), started on \(x\), halts within \(t\) steps with \(w\) on its output tape, then it also halts within \(t'\) steps with \(w\) on its output tape.
Let \(M\) be a finite Turing machine over an alphabet \(\Sigma \) and let \(x\) be a word over \(\Sigma \). If \(M\), started on \(x\), halts within \(t\) steps with \(w\) on its output tape and also halts within \(t'\) steps with \(w'\) on its output tape, then \(w = w'\). In other words, a machine has at most one completed output on a given input.
Let \(M\) be a finite Turing machine over an alphabet \(\Sigma \), let \(x\) and \(w\) be words over \(\Sigma \), and let \(t \in \mathbb {N}\). Then \(M\) computes the output \(w\) from the input \(x\) within \(t\) steps if and only if the configuration reached after running \(M\) for \(t\) steps from its initial configuration on \(x\) is halted and carries exactly the word \(w\) on its output tape.
Given three machines \(D\), \(M_1\), \(M_2\) over the Boolean alphabet with \(k_D\), \(k_1\), \(k_2\) work tapes respectively, the machine with \(k_D + k_1 + k_2\) work tapes whose finite control simulates \(D\) on the first tape block with \(D\)’s emissions suppressed and the first emitted symbol captured in a state register. After the simulated \(D\) halts, a control phase rewinds the input head to its initial position and then transfers control to a disjoint copy of the branch union of \(M_1\) and \(M_2\) selected by the register bit, running on the fresh remaining tape blocks; when the register is empty at dispatch the machine halts.
Let \(D\), \(M_1\), \(M_2\) be machines over the Boolean alphabet, let \(x\) be a Boolean word, and let \(b\) be a Boolean such that \(D\) halts on input \(x\) with the singleton output consisting of \(b\) alone (within some number of steps). Then there exist a time \(t\) and residual contents and head positions for the controller’s tape block such that the \(t\)-step run of the conditional-dispatch machine from its initial configuration on \(x\) is exactly the right-block embedding of the initial configuration on \(x\) of the branch machine selected by \(b\) (the branch union of \(M_1\) and \(M_2\)), alongside those residual controller tapes.
For a fixed Boolean word \(w\), the machine over the Boolean alphabet with no work tapes whose states are \(\{ 0, \dots , \left\lvert w\right\rvert \} \), forming the emission chain for \(w\): state \(i \lt \left\lvert w\right\rvert \), regardless of the symbols scanned, appends the \(i\)-th symbol of \(w\) (zero-indexed) to the output and passes to state \(i + 1\), and state \(\left\lvert w\right\rvert \) halts.
For a movement \(m \in \{ -1, 0, +1\} \) and an optional successor state \(q\), the one-step action of a machine with \(k\) work tapes over the Boolean alphabet that moves the input head by \(m\) (subject to clamping at the two boundary blanks), writes nothing on any work tape, moves no work head, emits no output, and passes to \(q\), halting when \(q\) is absent.
Let \(c\) be a configuration of a machine with \(k\) work tapes over the Boolean alphabet on input \(x\), let \(m \in \{ -1, 0, +1\} \) be a movement, and let \(q\) be an optional state. Applying the control action for \(m\) and \(q\) to \(c\) yields the configuration that agrees with \(c\) on all work-tape contents, work-head positions, and output, whose state is \(q\), and whose input head is that of \(c\) moved by \(m\) with clamping at the two boundary blanks.
Given machines \(D\), \(M_1\), \(M_2\) over the Boolean alphabet and a configuration \(c\) of \(D\) on an input \(x\), the configuration of the conditional-dispatch machine for \(D\), \(M_1\), \(M_2\) on \(x\) whose control state pairs the state of \(c\) with the first symbol (if any) of \(c\)’s output stored in the register, whose input head agrees with that of \(c\), whose first tape block carries \(c\)’s work tapes while all branch tapes are blank with heads at the origin, and whose real output is empty.
Let \(D\), \(M_1\), \(M_2\) be machines over the Boolean alphabet, let \(c\) be a configuration of \(D\) on an input \(x\), and let \(t\) be a number of steps such that the run of \(D\) from \(c\) has not halted at any time \(s \lt t\). Then the \(t\)-step run of the conditional-dispatch machine from the embedded controller configuration of \(c\) equals the embedded controller configuration of the \(t\)-step run of \(D\) from \(c\).
Let \(D\), \(M_1\), \(M_2\) be machines over the Boolean alphabet and let \(c\) be a configuration of \(D\) on an input \(x\) whose state is not the halting state. Then one step of the conditional-dispatch machine from the embedded controller configuration of \(c\) equals the embedded controller configuration of one step of \(D\) from \(c\); in particular the register tracks the first symbol of the simulated output, with no hypothesis on \(c\)’s output.
Let \(M\) be a finite multi-tape Turing machine over the binary alphabet, let \(L\) be a language over \(\{ 0,1\} \) (a set of finite binary strings), and let \(T : \mathbb {N} \to \mathbb {N}\) be a time bound. We say that \(M\) decides \(L\) within time \(T\) if on every input string \(x\) the machine halts within \(T(|x|)\) steps and produces the single-symbol output \(1\) when \(x \in L\) and \(0\) when \(x \notin L\).
Let \(\Gamma \) be an alphabet, let \(M\) be a deterministic finite Turing machine over \(\Gamma \), let \(e : \{ 0,1\} \hookrightarrow \Gamma \) be an injective embedding of the binary symbols into \(\Gamma \), let \(L \subseteq \{ 0,1\} ^{*}\) be a language over the binary alphabet, and let \(T : \mathbb {N} \to \mathbb {N}\) be a time bound. We say that \(M\) decides \(L\) via \(e\) within time \(T\) if for every binary string \(x\), when \(M\) is run on the input obtained by applying \(e\) to each symbol of \(x\), it halts within \(T(|x|)\) steps with output the one-symbol string \(e(b)\), where \(b \in \{ 0,1\} \) is the indicator of whether \(x \in L\).
For every partial tape selection, control embedding, default tapes and heads, source action \(a\), and source configuration \(c\): applying the relocated action to the relocated configuration of \(c\) equals the relocated configuration of \(a\) applied to \(c\), including the write, head motion, and emission; unselected tapes and heads are fixed.
For every mode, input, argument and capture words, time \(T\), and source tape \(i\): relocating the interval cleaner’s blank return endpoint into the cleanup frame whose first \(i\) banks are blank yields the cleanup frame whose first \(i + 1\) banks are blank, in the cleaner-return control state for bank \(i\).
For every mode, input, argument and capture words, time \(T\), and source tape \(i\): relocating the interval cleaner’s entry configuration for bank \(i\) (actual data, visited span, origin marker, actual head) into the cleanup frame whose first \(i\) banks are blank yields exactly that frame, with every other bank and both administrative tapes retained.
For every mode, input, argument and capture words, time \(T\), and source tape \(i\), there is a time \(t \le 6T + 8\) such that the clean-call controller, run for \(t\) steps from the cleanup frame with \(i\) cleaned banks in the cleaner-entry state for bank \(i\), reaches the cleanup frame with \(i + 1\) cleaned banks in the cleaner-entry state for bank \(i + 1\); the dispatch happens at the cleaner’s observed return.
For every mode, input, argument and capture words, time \(T\), and count \(j \le k\) (the number of source tapes), there is a time \(t \le j (6T + 8)\) such that the clean-call controller, run for \(t\) steps from the cleanup frame with no cleaned banks, reaches the cleanup frame with \(j\) cleaned banks, both administrative words preserved. The claim is a full configuration equality.
Let \(M\) be a machine, \(x\) an input, \(T\) a time, and \(i\) a work tape. Then there is a positive time \(t \le 6T + 7\) such that the interval cleaner, started on the actual tape \(i\) of the run of \(M\) at time \(T\) with its visited-span and origin markers, does not visit its return state before \(t\) and at time \(t\) is exactly the blank-endpoint return configuration with all heads at the origin.
Let \(d\) be a data tape and let an interval of width \(m\) start at \(\ell \). For every \(j \lt m\), running the interval cleaner for exactly \(j + 2\) steps from its initial state with all heads at \(\ell + j\) reaches its rightward-erasure state with all heads at \(\ell \), independently of blank holes in \(d\).
For every natural number \(n\), running the interval cleaner for exactly \(n + 1\) steps from its rewind state with blank data and marker tapes and all heads at \(n\) reaches the silent return state with all three tapes blank and all heads at the origin, the origin marker being erased by the final transition.
Let \(d\) be a data tape blank outside an interval of width \(m\) starting at \(\ell \), with \(\ell \le 0 \lt \ell + m\), and let the heads start at \(\ell + j\) for some \(j \lt m\). Then there is a positive time \(t \le 3m + 4\) after which the interval cleaner, started in its initial state on \(d\), the interval marker, and the origin marker, is in its silent return state with all three tapes blank and all heads at the origin.
Let \(d\) be a data tape and let an interval of width \(m\) start at \(\ell \). For every \(j \le m\), running the interval cleaner for exactly \(j\) steps from its rightward-erasure state with all heads at \(\ell \) erases the first \(j\) cells of both the data and the marker, moves all heads to \(\ell + j\), and leaves the origin marker intact.
For every data tape \(d\), integer \(\ell \), and count \(j\), blanking the cell \(\ell + j\) of the tape obtained from \(d\) by erasing the first \(j\) interval cells yields the tape obtained from \(d\) by erasing the first \(j + 1\) interval cells.
Let \(M\) be a machine with \(k\) work tapes computing output \(c\) on argument \(a\) within time \(T\). For every mode and input word, there is a time \(t \le (24 + 13k)(T + |a| + |c| + 1)\) such that the clean-call controller, run for \(t\) steps from its canonical entry seam on \(a\), reaches the canonical exit seam carrying \(c\) in install mode and \(a\) in emit mode, with output \(c\) exactly in emit mode.
Let \(M\) compute output \(o\) on argument \(s\) within time \(T\). Then there is a time \(t \le T\) and a boundary tag such that the run of \(M\) from its initial configuration on \(s\) is halted at \(t\) with output \(o\), the capture-wrapped evaluator started on the prepared candidate does not visit its return state before \(t\), and at time \(t\) it is exactly the evaluator configuration of that halted run.
For every physical input \(w\) and argument word \(s\), the evaluator configuration built from the initial configuration of \(M\) on \(s\), with initial tag and physical input head at one, equals the canonical seam configuration of the capture-wrapped evaluator whose first work tape holds \(s\), all other work tapes blank, all heads at their origins, and output empty.
Let \(M\) be a bundled finite machine, \(c\) a configuration of \(M\) on a simulated argument carrying a valid virtual boundary tag, and \(t\) a time before which the run of \(M\) from \(c\) stays live. Then there is a valid tag for the endpoint such that running the capture-wrapped evaluator for \(t\) steps from the evaluator configuration of \(c\) equals the evaluator configuration of the run of \(M\) from \(c\) at time \(t\).
For every mode, input, argument word \(a\), capture word \(c\), and time \(T\): running the clean-call controller for exactly \(|a| + 3|c| + 6\) steps from the cleanup frame in which all banks are clean (in the last cleaner-dispatch state) reaches the canonical exit seam carrying \(c\) in install mode and \(a\) in emit mode, with output \(c\) exactly in emit mode, every scratch tape blank, and every head at its origin.
Let the finalizer run in a fixed mode on an argument word \(a\) and a capture word \(c\), and let \(p\) be a prefix of \(a\). Starting from its argument-scan state with the argument head at cell \(|p| - 1\) – the argument tape holding \(a\) in emit mode and only \(p\) in install mode – after exactly \(|p| + 1\) steps it is in the capture-rewind state with the argument tape holding \(a\) in emit mode and empty in install mode, the argument head at the origin, and the capture head one cell left of \(|c|\).
Let the finalizer run in a fixed mode with a remaining captured word \(p\). Starting from its erasure state with the capture head at cell \(|p| - 1\) (and, in install mode, the argument head moving along with it), after exactly \(|p| + 1\) steps it is in its final silent state with the capture tape completely blank and both heads at the origin, the output untouched.
For every mode, input, argument word \(a\), and capture word \(c\): relocating the finalizer’s final configuration through the pair layout equals the canonical seam in the exit control state whose first work tape holds \(c\) in install mode and \(a\) in emit mode, all other tapes blank, all heads at their origins, with output \(c\) in emit mode and empty otherwise.
For every mode, input, argument word \(a\), capture word \(c\), and time \(T\): relocating the finalizer’s start configuration (argument head at \(|a|\), capture head at \(|c|\), empty output) through the pair layout equals the cleanup frame in which all source banks are blank, in the finalizer-entry control state.
Let the finalizer run in a fixed mode with capture word \(c\), and let \(p\) be a prefix of \(c\). Starting from its capture-rewind state with the capture head at cell \(|p| - 1\), after exactly \(|p| + 1\) steps it is in the transfer state with the capture head at the origin, the capture word intact, and the argument head still at the origin.
Let the finalizer run in a fixed mode on an argument word \(a\) and a capture word \(c\), entering at its start state with the argument head at \(|a|\) and the capture head at \(|c|\) and empty output. After exactly \(|a| + 3|c| + 5\) steps it is in its final silent state with both heads at the origin and the capture tape blank; the argument tape holds \(c\) in install mode and \(a\) in emit mode, and the output is \(c\) in emit mode and empty in install mode. Empty words are included.
Let the finalizer run on an argument word \(a\) and capture word \(c\), and write \(c\) as a processed prefix \(p\) followed by a remainder \(r\). Starting from its transfer state with the capture head at \(|p|\) – having already transferred \(p\) to the argument tape (install mode) or the output (emit mode) – after exactly \(|r| + 1\) steps it is in the erasure state with the full capture transferred: the argument tape holds \(c\) in install mode and \(a\) in emit mode, the output holds \(c\) exactly in emit mode, and the capture word itself is still intact.
Let \(M\) be a machine with \(k\) work tapes computing output \(c\) on argument \(a\) within time \(T\). For every mode and input word, there is a positive time \(t \le (24 + 13k)(T + |a| + |c| + 1)\) such that the clean-call controller, started at its canonical entry seam on \(a\), does not visit the exit state before \(t\) and at time \(t\) is exactly the canonical exit seam carrying \(c\) in install mode and \(a\) in emit mode, with output \(c\) exactly in emit mode.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet, let a set of stop control states be absorbing (a step from any configuration in a stop state changes nothing), and suppose the run from a configuration \(c\) reaches a configuration \(d\) in a stop state at time \(T\). Then there is a time \(t \le T\) such that no stop state is visited strictly before \(t\) and the run from \(c\) at time \(t\) equals \(d\).
For every assignment of per-slot values and every source tape \(i\), the clean-call layout evaluated at the physical index of component \(r\) of bank \(i\) gives the data value, the visited value, or the origin value of tape \(i\) according to \(r\).
Let the source run on argument \(a\) have output \(c\) at time \(T\). The identity relocation of the evaluator configuration of the halted right-normalized tracked endpoint equals the cleanup frame with zero cleaned banks: complete source data, visited spans, origin markers, captured result \(c\), and the argument head at its right boundary.
For every mode, input word, and argument word \(a\): the identity relocation of the prepared evaluator’s entry configuration on \(a\) equals the clean-call controller’s canonical entry seam, whose first work tape holds \(a\) with all other tapes blank, heads at their origins, and empty output.
Let \(M\) compute output \(c\) on argument \(a\) within time \(T\). For every mode and input word, there is a time \(t \le 2T + |a| + 4\) such that the clean-call controller, run for \(t\) steps from its canonical entry seam on \(a\), reaches the cleanup frame with zero cleaned banks in the first cleaner-entry state, with no scan of the native input.
Let \(M\) compute output \(o\) on argument \(s\) within time \(T\), and consider the capture- wrapped evaluator of the right-boundary wrapper of the tracked evaluator of \(M\), started on the prepared candidate seam for \(s\). Then there is a positive time \(t \le 1 + 2T + |s| + 2\) such that the evaluator does not visit its return state before \(t\) and at time \(t\) is exactly the evaluator configuration of the halted, right-normalized tracked endpoint: full trace banks, complete capture buffer, and candidate head at the right boundary.
Let a source machine, a host machine, a tape index map with a partial inverse selection, and a control embedding be given such that on all good source states the host transition equals the relocated source transition. If every live state visited by the source run from a configuration \(c\) before time \(t\) is good, then the host run for \(t\) steps from the relocated configuration of \(c\) equals the relocated configuration of the source run at time \(t\).
Let \(M\) compute output \(o\) on input \(x\) within time \(T\). Then the right-boundary wrapper of \(M\) computes \(o\) on \(x\) within time \(T + |x| + 2\), and after exactly that many steps its input head is at position \(|x| + 1\).
Let \(M\) be a machine and \(x\) an input such that the run of \(M\) from its initial configuration is halted at time \(T\). Then the right-boundary wrapper of \(M\), run for \(T + |x| + 2\) steps from its initial configuration on \(x\), equals the halted scan configuration built on the run of \(M\) at time \(T\), with input head at position \(|x| + 1\).
Let \(c\) be a halted configuration of the source machine on an input \(x\). Then there is a time \(t \le |x| + 2\) such that the right-boundary wrapper, run for \(t\) steps from the embedding of \(c\), is the halted scan configuration at position \(|x|\), with the work tapes and output of \(c\) unchanged.
Let \(c\) be a configuration of the source machine and \(t\) a time before which the source run from \(c\) stays live. Then the right-boundary wrapper run for \(t\) steps from the embedding of \(c\) equals the embedding of the source run at time \(t\).
Let \(c\) be a completed source configuration on an input \(x\) and let \(j\) and \(r\) satisfy \(j + r = |x|\). Running the right-boundary wrapper for exactly \(r + 1\) steps from the live scan configuration at position \(j\) halts it silently at the right-boundary scan configuration at position \(|x|\), never confusing a blank work cell with the input end.
Let \(c\) be a live configuration of the source machine. One step of the right-boundary wrapper from the embedding of \(c\) equals the embedding of the one-step successor of \(c\), including any halting emission; only the successor control encoding differs.
Let \(c\) be a live configuration of the source machine with visited bounds \(\ell \) and \(h\). One step of the tracked evaluator from the completed tracked configuration of \(c\) equals the intermediate stamp configuration of the one-step successor of \(c\) with the same bounds: the data bank performs the source write and emission, and all three per- tape heads move by the same source action.
Let \(M\) be a machine, \(x\) an input, \(T\) a time, and \(i\) a work tape. Then there is a positive time \(t \le 6T + 7\) after which the interval cleaner, started on the actual tape \(i\) of the run of \(M\) at time \(T\), the closed-span marker of the visited interval, and the origin marker, with heads at the actual head position, reaches its return state with all three tapes blank and all heads at the origin.
For every machine \(M\), input \(x\), time \(t\), and work tape \(i\): the leftmost visited position is at most zero, the rightmost is at least zero, the current head at time \(t\) lies between them, and they lie within \([-t, t]\), so the visited interval has width at most \(2t + 1\).
For every machine \(M\) and input \(x\), one step of the tracked evaluator from its initial configuration on \(x\) equals the completed tracked configuration of the initial configuration of \(M\) on \(x\) with all visited bounds zero: both marker banks hold origin marks, the data bank is blank, and the output is empty.
For every machine \(M\), input \(x\), and time \(t\), running the tracked evaluator for \(1 + 2t\) steps from its initial configuration on \(x\) equals the completed tracked configuration of the run of \(M\) at time \(t\) with the true leftmost and rightmost visited bounds, including after the source has halted.
Let \(c\) be a source configuration with visited bounds \(\ell _i \le h_i\) such that each head of \(c\) lies within one cell of its span. One step of the tracked evaluator from the intermediate stamp configuration of \(c\) equals the completed tracked configuration of \(c\) with each span extended to include the current head, after which the stored successor control (possibly a halt) is dispatched.
For every machine \(M\), input \(x\), time \(t\), work tape \(i\), and cell \(z\) outside the visited interval of tape \(i\) at time \(t\), the cell \(z\) of tape \(i\) after running \(M\) for \(t\) steps from its initial configuration on \(x\) is blank.
For a partial selection assigning to each host tape an optional source tape, an embedding of source control states into host states, and a source action, the host action whose selected tapes receive the source operations, whose unselected tapes are stationary, and whose input move, emission, and successor state are those of the source action under the state embedding.
For a mode bit, input and argument and capture words, a time \(T\), a count \(j\), and a control state, the cleanup frame in which the first \(j\) source banks are completely blank with heads at the origin, while every remaining bank holds the exact halted trace of the source at time \(T\): its data tape, the closed-span marker of its visited interval, the origin marker, and its true head position.
For a partial tape selection, a control-state embedding, default host tapes and heads for the unselected slots, and a source configuration \(c\), the host configuration whose selected tapes and heads come from \(c\), whose unselected tapes and heads are the defaults, and whose control, input position, and output are those of \(c\) under the state embedding.
For an input word, a physical input-head position, a control state, three tape functions for data, interval marks, and the origin marker, and a common head position \(h\), the configuration of the three-tape interval cleaner whose tapes are as given, whose three work heads are all at \(h\), and whose output is empty.
For a data tape \(d\), an integer \(\ell \), and a count \(j\), the tape that is blank on the cells \(z\) with \(\ell \le z \lt \ell + j\) and agrees with \(d\) everywhere else: the contents after erasing the first \(j\) cells of the interval.
A bundled finite machine over the Boolean alphabet with three work tapes – data, an interval marker, and an origin marker – and four control states, that scans left to the marked interval’s left end, erases the data and the marker in one rightward pass, rewinds on the surviving origin marker, and finishes in a silent absorbing return state with all three heads at the origin. All three heads move together, and only observed tape symbols dispatch control.
For a configuration \(c\) of a machine \(M\) on a simulated argument, a virtual boundary tag, and a physical input-head position, the configuration of the capture-wrapped evaluator in which the preserved argument occupies tape zero, the simulated work bank of \(c\) follows it, the final tape captures every emission of \(c\), and the physical input head is fixed at the given position.
For a bundled finite machine \(M\), the evaluator obtained by wrapping the buffered virtual-input simulator of \(M\) (reading its argument from a designated work tape rather than the native input) with the capture transformation: every emission of the simulated run lands on a final capture tape, the completed state is a live self-looping return state, and no evaluated bit reaches the physical output.
For a mode bit, an input word, a physical input-head position, one of six control states, two tape functions, two head positions, and an output word, the two-tape finalizer configuration whose argument and capture tracks are the two given tapes with the given heads and whose control, input position, and output are as given.
A bundled finite machine over the Boolean alphabet with two work tapes (the argument and the captured result) and six control states that, entering at the known right blanks of both words, rewinds the argument – erasing it in install mode, preserving it in emit mode – rewinds the capture, transfers the captured word in order (writing it onto the argument tape in install mode, emitting it to the physical output in emit mode), then erases the capture and returns both heads to the origin, finishing in an absorbing silent state.
For a mode bit, an input word, an argument word \(a\), a capture word \(c\), a control state, per-tape data, visited, and origin banks, and per-tape head positions, the clean- call configuration whose argument track is the buffer tape of \(a\) with head at \(|a|\), whose capture track is the buffer tape of \(c\) with head at \(|c|\), whose three banks are as given with their common per-tape heads, with input head at one and empty output.
For a machine \(M\), an input \(x\), a time \(t\), and a work tape \(i\), the maximum of the origin and all head positions of tape \(i\) along the run of \(M\) from its initial configuration on \(x\) through time \(t\); it is the right end of the visited interval of tape \(i\).
A bundled finite machine over the Boolean alphabet with no work tapes and a single control state whose only transition is the stationary silent self-loop. It supplies the empty inactive bank of the virtual-input wrapper; its own transition is never entered in that use.
For per-slot values – an argument value, a capture value, and data, visited, and origin values for each source tape – the family over all physical call-layout tapes whose order is: argument, data bank, visited bank, origin bank, capture.
For a machine \(M\), an input \(x\), a time \(t\), and a work tape \(i\), the minimum of the origin and all head positions of tape \(i\) along the run of \(M\) from its initial configuration on \(x\) through time \(t\); it is the left end of the visited interval of tape \(i\).
For a configuration \(c\) of the source machine, the configuration of the right-boundary wrapper with the same tapes, heads, and output, whose control is the embedded live state of \(c\) when \(c\) is live and the administrative scan-entry state when \(c\) has halted.
For a completed source configuration \(c\), an optional administrative control bit, and a position \(j\) at most the input length, the wrapper configuration whose work tapes, heads, and output are those of \(c\) verbatim and whose input head is at position \(j + 1\); only the input position moves during the scan.
For a bundled finite machine \(M\), the machine that follows \(M\) exactly while it is live and, after the halting transition of \(M\), moves the input head rightwards to the right boundary blank and then halts silently; the completed source output and work tapes are untouched by the scan.
For a bundled finite machine \(M\), the capture-wrapped evaluator of the right-boundary wrapper of the tracked evaluator of \(M\): one genuine argument tape, three tracked banks, and one capture tape, even when \(M\) has no work tapes.
For a bundled finite machine \(M\) and a mode bit, the controller machine that evaluates \(M\) on a tape-resident argument through the prepared tracked source and captures the result at its observed return, then sequentially runs the interval cleaner on each tracked bank, each dispatched at its observed return, and finally installs the captured result on the argument tape (install mode) or forwards it to the physical output while preserving the argument (emit mode), erasing the capture. No deadline or source running time occurs in its transition table.
For a configuration \(c\) of the source machine and per-tape visited bounds \(\ell \) and \(h\), the tracked-evaluator configuration whose data bank carries \(c\), whose visited bank holds the closed-span markers \([\ell _i, h_i]\), whose origin bank marks each origin, and whose three per-tape heads agree with the heads of \(c\).
For a source configuration \(c\) and visited bounds, the tracked-evaluator configuration that differs from the completed tracked configuration only in its control state, which is the stamping state storing the successor control of \(c\): the data, input, heads, and emitted output already reflect the performed action while the visited markers are still the old ones.
For a bundled finite machine \(M\), the machine with three parallel tape banks that simulates \(M\) at two native steps per source step: the first microstep performs the source action on the data bank while moving the marker heads identically, and the second stamps the newly reached cells in the visited bank before dispatching, possibly halting. A single initialization step marks each origin in both marker banks. The physical output is the source output.
Let \(M\) be a multi-tape Turing machine with \(k\) work tapes over the Boolean alphabet running on some input \(x\), let \(w\) be a Boolean word, and let \(e\) be a map from \(\{ 0,\dots ,\left\lvert w\right\rvert \} \) into the states of \(M\) such that from each state \(e(i)\), regardless of the symbols read, the transition of \(M\) is the emission action for \(w\) at index \(i\) (append the \(i\)-th symbol of \(w\) and pass to \(e(i+1)\) while \(i \lt \left\lvert w\right\rvert \), halt at \(i = \left\lvert w\right\rvert \), never moving the input head or touching a work tape). If a configuration \(c\) is in state \(e(0)\), then after exactly \(\left\lvert w\right\rvert + 1\) steps from \(c\) the machine has halted and its output is the output of \(c\) followed by all of \(w\).
Let \(M\) be a multi-tape Turing machine with \(k\) work tapes over the Boolean alphabet running on some input \(x\), let \(w\) be a Boolean word, and let \(e\) be a map from \(\{ 0,\dots ,\left\lvert w\right\rvert \} \) into the states of \(M\) such that from each state \(e(i)\), regardless of the symbols read, the transition of \(M\) is the emission action for \(w\) at index \(i\): while \(i \lt \left\lvert w\right\rvert \) it appends the \(i\)-th symbol of \(w\) to the output and passes to \(e(i+1)\), and at \(i = \left\lvert w\right\rvert \) it halts, in both cases leaving the input head and all work tapes fixed. If a configuration \(c\) is in state \(e(0)\), then for every \(t \le \left\lvert w\right\rvert \), after \(t\) steps from \(c\) the machine is in state \(e(t)\) and its output is the output of \(c\) followed by the first \(t\) symbols of \(w\).
Given a Boolean word \(w\), a map \(e\) from \(\{ 0,\dots ,\left\lvert w\right\rvert \} \) into the states of a machine with \(k\) work tapes over the Boolean alphabet, and an index \(i \in \{ 0,\dots ,\left\lvert w\right\rvert \} \), the one-step action that, when \(i \lt \left\lvert w\right\rvert \), appends the \(i\)-th symbol of \(w\) (zero-indexed) to the output and passes to state \(e(i+1)\), and otherwise emits nothing and halts. In both cases the input head and all work tapes are left untouched.
For a candidate \(s\) with \(\left\lvert s\right\rvert \le \left\lvert w\right\rvert \), the binary representation of \(f(\left\lvert s\right\rvert )\) equals the binary representation of the length of the suffix of \(w\) after position \(\left\lvert s\right\rvert \) if and only if \(\left\lvert s\right\rvert + f(\left\lvert s\right\rvert ) = \left\lvert w\right\rvert \); this includes the case of two empty words.
For each work tape of a source machine run for \(T\) steps, there is a time \(t\) with \(0 \lt t \le 6T + 7\) such that the interval cleaner, started on that tape’s data with its tracked span and origin marker, never visits its return state before time \(t\) and at time \(t\) has reached the all-blank return configuration.
Let a marked interval have left endpoint \(m\) and width \(W\), and let \(j \lt W\). Then \(j+2\) steps of the cleaner, started in its initial scan state at offset \(j\) from the left end, yield its clearing state at the left end \(m\), regardless of blank holes in the data; the data, interval marker, and origin marker are unchanged.
For every \(n\), if the data and interval-marker tapes are already blank and the cleaner is in its rewind state at displacement \(n\), then exactly \(n+1\) steps return all heads to the origin and erase the origin marker, reaching the return state with all three tapes blank.
Let a marked interval of width \(W\) contain the origin (\(\mathrm{left} \le 0 \lt \mathrm{left} + W\)), let the data vanish outside it, and start the cleaner anywhere inside it. Then there is a time \(t\) with \(0 \lt t \le 3W + 4\) after which the cleaner is in its return state with all three tapes blank and all heads at the origin.
Let a marked interval have left endpoint \(m\) and width \(W\), and let \(j \le W\). Then \(j\) steps of the cleaner from its clearing state at \(m\) yield the configuration in which the first \(j\) cells of both the data and the interval marker are cleared and all heads are at offset \(j\), while the origin marker is intact.
For words \(u, v\), there is a time \(t\) with \(0 \lt t \le 2(\max (\left\lvert u\right\rvert ,\left\lvert v\right\rvert ) + 1)\) such that the comparator never visits either Boolean return state before time \(t\) and at time \(t\) sits in its return configuration with the register recording whether \(u = v\) and both heads at zero.
From the comparator’s rewind state with both heads at displacement \(j-1\), for any \(j \le \max (\left\lvert u\right\rvert , \left\lvert v\right\rvert )\), exactly \(j+1\) steps restore both heads to displacement \(0\) and enter the return state, preserving the comparison register.
For all words \(u\) and \(v\), whole-word comparison of \(u\) and \(v\) completes silently in exactly \(2(\max (\left\lvert u\right\rvert , \left\lvert v\right\rvert ) + 1)\) steps, ending in the return state with the register recording whether \(u = v\), both word tapes unchanged, and both heads back at zero.
For every \(j \le \max (\left\lvert u\right\rvert , \left\lvert v\right\rvert )\), after \(j\) forward steps of the comparator the equality register records exactly whether the length-\(j\) prefixes of \(u\) and \(v\) agree, with both heads at \(j\); this includes mismatches caused by unequal lengths.
Let \(M\), on the candidate \(s\), halt within \(T\) steps with output \(o\). Then there is a time \(t \le T\) at which \(M\)’s run on \(s\) has first halted, with output \(o\) and a valid virtual tag, such that the candidate evaluator never visits its return state before time \(t\) and at time \(t\) sits in the evaluator configuration of \(M\)’s halted configuration, its capture tape holding exactly \(o\).
The evaluator configuration of the simulated machine’s initial configuration on a candidate \(s\) — candidate buffered, source work bank blank, capture tape empty, all heads at zero — equals the canonical anchored configuration whose stored words carry \(s\), at the evaluator’s initial virtual state.
Let the simulated machine, started from a configuration \(c\) over the candidate with a valid virtual-input tag, stay live for all times before \(t\). Then \(t\) steps of the candidate evaluator from the evaluator configuration of \(c\) land exactly on the evaluator configuration of the simulated machine’s configuration at time \(t\), with a valid tag.
Let a class of control states be absorbing (any configuration whose state lies in the class is a fixed point of the step function), and suppose the run from \(c\) reaches a configuration \(d\) with state in the class after \(T\) steps. Then there is a time \(t \le T\) with no earlier visit to the class such that the configuration at time \(t\) is already \(d\).
Let \(M\), on the candidate \(s\), halt within \(T\) steps with output \(o\), and let the composite evaluator capture the tracked, right-normalized simulation of \(M\). Then there is a time \(t\) with \(0 \lt t \le 1 + 2T + \left\lvert s\right\rvert + 2\) such that the composite never visits its return state before \(t\) and at time \(t\) sits in the evaluator configuration carrying the full tracked trace banks, the complete captured output, the candidate head at the right boundary, nothing on the physical output, and the physical input head fixed at position one.
If a machine halts on input \(x\) with output \(o\) within \(T\) steps, then its right-boundary wrapper halts on \(x\) with output \(o\) within \(T + \left\lvert x\right\rvert + 2\) steps, and at that time its input head sits at position \(\left\lvert x\right\rvert + 1\).
If a source machine has halted on input \(x\) by time \(T\), then \(T + \left\lvert x\right\rvert + 2\) steps of its right-boundary wrapper from the initial configuration reach the halted scan configuration carrying the complete source configuration at time \(T\) — work tapes included — with the input head at the right boundary.
For a halted source configuration \(c\), there is a time \(t \le \left\lvert x\right\rvert + 2\) after which the right-boundary wrapper, started from the wrapper configuration of \(c\), has reached the halted scan configuration at the right boundary; neither the mandatory initial positive move nor the scan changes the completed work tapes or output.
Let \(x\) be an input and let \(c\) be a source configuration on \(x\). For all \(j\) and \(r\) with \(j + r = \left\lvert x\right\rvert \), exactly \(r + 1\) silent steps of the right-boundary wrapper from the scan configuration of \(c\) at position \(j\) reach the halted scan configuration of \(c\) at position \(\left\lvert x\right\rvert \); the scan never mistakes a blank work cell for the end of the input.
For a live source configuration \(c\), one step of the right-boundary wrapper from the wrapper configuration of \(c\) equals the wrapper configuration of the source configuration after one step, including any halting emission; only the control-state encoding differs.
For words \(u, v\) and an index \(j\), the length-\((j+1)\) prefixes of \(u\) and \(v\) agree if and only if their length-\(j\) prefixes agree and their optional \(j\)-th entries agree; in particular a missing bit differs from a present bit.
For a live source configuration \(c\), one step of the tracked evaluator from the tracked configuration of \(c\) yields the intermediate stamp configuration of the source configuration after one source step: the data simulation is exact (including any halting emission), and both marker heads move with the data head without writing.
For each work tape of a source machine run for \(T\) steps, there is a time \(t\) with \(0 \lt t \le 6T + 7\) after which the interval cleaner, started on that tape’s data together with its tracked visited span and origin marker at the tape’s current head position, reaches its return state with all three tapes blank and all heads at the origin.
For every time \(t\) and work tape \(i\), the visited interval \([\mathrm{lo}, \mathrm{hi}]\) of the run contains \(0\) and the current head position, and satisfies \(-t \le \mathrm{lo}\) and \(\mathrm{hi} \le t\); in particular its width is at most \(2t + 1\).
For every machine \(M\) and input \(x\), one step of the tracked evaluator for \(M\) from its initial configuration on \(x\) yields the tracked configuration of the initial configuration of \(M\) on \(x\) with every visited span equal to \([0, 0]\): both origin markers are installed, the source work tapes are blank, and the output is empty.
For every time \(t\), the tracked evaluator’s configuration after \(1 + 2t\) native steps is exactly the tracked configuration of the source machine’s configuration after \(t\) steps, with visited intervals given by the leftmost and rightmost visited positions through time \(t\); this holds also after the source has halted.
Suppose each visited span \([\mathrm{lo}_i, \mathrm{hi}_i]\) is nonempty and each current source head lies within one cell of it. Then one step of the tracked evaluator from the intermediate stamp configuration marks every current head cell, producing the tracked configuration whose spans are extended to \([\min (\mathrm{lo}_i, h_i), \max (\mathrm{hi}_i, h_i)]\), and only then dispatches the stored successor state.
Let a machine \(E\) compute the binary width word \(s \mapsto \mathrm{bits}(f(\left\lvert s\right\rvert ))\) within a monotone time bound \(T_E\). Then for any candidate \(s\) of length at most \(\left\lvert w\right\rvert +1\): \(E\) halts on \(s\) with that output within \(T_E(\left\lvert s\right\rvert )\) steps, the output’s length is at most \(T_E(\left\lvert s\right\rvert )\), and \(T_E(\left\lvert s\right\rvert ) \le T_E(\left\lvert w\right\rvert +1)\).
Let a machine compute the binary width word within a monotone bound \(T_E\), and let the candidate \(s\) have length at most \(\left\lvert w\right\rvert + 1\). Then the composite evaluator has a first return at some positive time \(t \le 3(T_E(\left\lvert w\right\rvert +1) + \left\lvert w\right\rvert + 2)\), at which it retains its tracked work banks and the exact captured width word.
After a source machine has run for \(T\) steps, exactly \(6T + 7\) steps of the simultaneous cleaner bank, started on the assembly of all tracked tapes with their spans and origin markers, clear the entire bank: all data, marker, and origin tapes become blank with every work head at zero, while the physical input head and output are unchanged; this includes a source with no work tapes.
After a source machine has run for \(T\) steps, there is a time \(t \le 6T + 7\) such that the simultaneous cleaner bank, started on the assembly of all tracked tapes, never shows the all-returned state vector before time \(t\) and at time \(t\) has reached the fully blank assembled endpoint; a bank over zero tapes may return at time zero.
Selecting component \(i\) of the assembled bank configuration recovers exactly the \(i\)-th cleaner’s transition on its own configuration, including all three independently positioned work heads; a halted component yields the stationary action.
Given \(k\) interval-cleaner configurations on a physical input \(w\) and an input-head position \(p\), the assembled bank configuration is the configuration of the simultaneous cleaner bank whose control state is the vector of the \(k\) cleaner states, whose data, visited-marker, and origin-marker banks carry the corresponding tapes and head positions of the \(k\) cleaners, whose input head is at \(p\), and whose physical output is empty.
The action performed by one component of the simultaneous cleanup: a live component takes its interval cleaner’s transition on its own triple of symbols, and a stopped component performs the stationary silent action; the native input is ignored.
A four-state machine with three work tapes — data, interval marker, and origin marker — that clears a finite marked work interval natively: it scans left to the marked interval’s left end, clears data and marker in lockstep moving right, rewinds using the origin marker, and erases that marker at the origin, stalling silently in its return state. Blank holes in the data are permitted; all heads stay aligned.
Given a physical input \(w\), words \(u\) and \(v\), an input position \(p\), a control phase \(q\), a register bit \(b\), and a displacement \(h\), the comparator configuration is the configuration of the two-word comparator in state \((q, b)\) with input head at \(p\), with \(u\) on its first tape and \(v\) on its second, both heads at displacement \(h\), and empty physical output.
A machine with two read-only word tapes that scans both in lockstep through their common right blank, accumulating in its finite control the Boolean record of whether all cells so far agree, then rewinds both heads to zero; the phase never emits, and empty words follow the same positive-time path.
A configuration of the candidate evaluator in which the preserved candidate occupies the first work tape, the simulated machine’s work bank follows it, and the final tape captures every emission of the simulated machine; the physical input head stays at a fixed position and the physical output is empty.
Given a machine \(M\), the candidate evaluator obtained by wrapping the virtual-input phase of the buffered composition over \(M\) with the capture transformer: \(M\) runs on a buffered candidate as virtual input, every emission lands on a final capture tape rather than the physical output, and completion passes to a live return state.
For a machine \(M\), an input \(x\), and a work tape, the rightmost visited position at time \(0\) is \(0\), and at time \(t+1\) it is the maximum of its value at time \(t\) and the position of that work head after \(t+1\) steps of \(M\) on \(x\).
For a machine \(M\), an input \(x\), and a work tape, the leftmost visited position at time \(0\) is \(0\), and at time \(t+1\) it is the minimum of its value at time \(t\) and the position of that work head after \(t+1\) steps of \(M\) on \(x\).
If an anchor-free run of length \(t\) ends on a canonical word seam in a non-anchor state whose transition is the silent dispatch to another non-anchor state, then the run of length \(t+1\) ends on the same word seam in the new state and is still anchor-free.
Applying a relocated action to a relocated configuration equals relocating the result of applying the original action: the write, head motion, and emission are reproduced on the selected slots, and unselected tape contents and positions are fixed.
There is a time \(t \le 2\left\lvert s\right\rvert + \left\lvert w\right\rvert + 5\) after which the controller, from its advance seam carrying the candidate \(s\), has replaced \(s\) by one split-search step applied to \(s\) (appending a \(1\) exactly within the input range, and stalling otherwise) and reached the advance phase’s return seam; the trace avoids the anchor.
For installed result words \(u\) and \(v\), there is a time \(t \le 2(\max (\left\lvert u\right\rvert ,\left\lvert v\right\rvert ) + 1)\) after which the controller, from its comparison seam, records in finite control whether \(u = v\), changing no candidate, buffer, or module scratch cell; the trace avoids the anchor.
If the candidate’s length is at most \(\left\lvert w\right\rvert \), then exactly \(\left\lvert s\right\rvert + \left\lvert w\right\rvert + 3\) controller steps from the emission seam halt with output the self-delimiting pair encoding of the split of the native input \(w\) at position \(\left\lvert s\right\rvert \); the emitted bits come from the native input, not the candidate, unrelated banks stay blank, and the trace avoids the anchor.
For any verdict bit, there is a time \(t \le 2(\left\lvert u\right\rvert +1)\) after which the controller, from its first erasure seam, has blanked the installed width word \(u\) while retaining the verdict in control and preserving every other tape; the trace avoids the anchor.
For any verdict bit, there is a time \(t \le 2(\left\lvert v\right\rvert +1)\) after which the controller, from its second erasure seam, has blanked the installed length word \(v\) while retaining the verdict through the last cleanup phase; after this all scratch and administrative words are blank. The trace avoids the anchor.
Let a verdict bit be given, true only if the candidate’s length is at most \(\left\lvert w\right\rvert \). Then there is a time \(t \le 2\left\lvert u\right\rvert + 2\left\lvert v\right\rvert + 2\left\lvert s\right\rvert + \left\lvert w\right\rvert + 12\) after which the controller, from the first erasure seam, has erased both installed words and: on a true verdict, halted with output the pair encoding of the split of \(w\) at position \(\left\lvert s\right\rvert \); on a false verdict, returned to the canonical anchored seam carrying one split-search step applied to \(s\). No strictly interior configuration is at the anchor.
If the suffix-length machine, started on its canonical seam carrying the prepared suffix, reaches its exit seam with result word \(z\) in \(t \gt 0\) steps with no interior exit-state visit, then \(t\) controller steps from the length-start seam install \(z\) in the length bank while preserving the candidate and the already installed width word; the trace avoids the anchor.
For any input \(w\) and candidate \(s\), there is a time \(t \le 5(\left\lvert s\right\rvert + \left\lvert w\right\rvert + 3)\) such that the controller, from its preparation seam carrying \(s\), reaches the seam holding the candidate, its copy, and the native suffix of \(w\) after position \(\left\lvert s\right\rvert \), with the past-the-end flag recording \(\left\lvert w\right\rvert \lt \left\lvert s\right\rvert \); the trace avoids the anchor.
Suppose the width call on the candidate \(s\) returns the binary word of \(f(\left\lvert s\right\rvert )\) in \(a \gt 0\) steps and the suffix-length call returns the binary word of the suffix length in \(b \gt 0\) steps, each with no interior exit-state visit. Then there is a time \(t\) with \(0 \lt t \le a + b + 10(\left\lvert s\right\rvert + \left\lvert w\right\rvert + \ell _u + \ell _v + 10)\), where \(\ell _u, \ell _v\) are the two binary words’ lengths, such that the round from the canonical anchored seam carrying \(s\) never revisits the anchor strictly between \(0\) and \(t\), and at time \(t\): if \(\left\lvert s\right\rvert + f(\left\lvert s\right\rvert ) = \left\lvert w\right\rvert \) the controller has halted with output the pair encoding of the split of \(w\) at \(\left\lvert s\right\rvert \), and otherwise it is back at the canonical anchored seam carrying the stepped candidate.
Suppose the width call completes in \(a \gt 0\) steps with result \(u\) and the suffix-length call completes in \(b \gt 0\) steps with result \(v\), each with no interior exit-state visit. Then there is a time \(t \le a + b + 2(\max (\left\lvert u\right\rvert ,\left\lvert v\right\rvert ) + 1) + 4\) after which the controller, from its completed-preparation seam, has run both calls and the whole-word comparison and sits at the first erasure seam carrying the verdict of \(u = v\); both calls are charged before any relation between their results is assumed, and the trace avoids the anchor.
If the width machine, started on its canonical seam carrying the candidate copy, reaches its exit seam with result word \(u\) in \(t \gt 0\) steps with no interior exit-state visit, then \(t\) controller steps from the width-start seam install \(u\) in the width bank while retaining the original candidate and prepared suffix; the trace avoids the anchor.
Let a host enter a relocated module through a dedicated entry control state that performs the module entry state’s action unconditionally, and agree with the relocated module transitions away from the module’s exit state. If the module runs from its entry configuration to \(d\) in \(t \gt 0\) steps with no interior visit to the exit state, then \(t\) host steps from the entry configuration land on the relocation of \(d\), with the whole trace avoiding the host anchor; this covers modules whose entry and exit states coincide.
Let a machine \(E\) compute the binary representation of \(f\) applied to the input length within a monotone time bound \(T_E\). Then there exist a machine \(M\) and a constant \(c\) computing, within \(c\, (n+1)(T_E(n+1) + n + 2)\) steps, the function sending \(w\) to the pair encoding of the split of \(w\) at the least \(i\) with \(i + f(i) = \left\lvert w\right\rvert \), and to the empty word when no such \(i\) exists.
If a control state’s transition is the unconditional silent dispatch to a state \(r\), then one step from the canonical word-seam configuration in that state yields the canonical word-seam configuration in state \(r\) with every other component unchanged.
From the eraser’s backward state with the word \(u\) and head at displacement \(\left\lvert u\right\rvert - 1\), exactly \(\left\lvert u\right\rvert + 1\) steps erase every cell of \(u\) right to left and restore the head to the origin, reaching the return state with a blank tape.
For any word \(u\), there is a time \(t\) with \(0 \lt t \le 2(\left\lvert u\right\rvert +1)\) such that the eraser never visits its return state before \(t\) and at time \(t\) has reached the blank-tape return configuration; the return is positive even for the empty word.
Relocating a canonical state-word configuration of the width-call machine into the host layout — with the candidate and suffix words preserved on their ambient tapes — yields exactly the host’s canonical configuration on the full word layout with the width word installed.
For a host slot, a host word layout, a physical input \(w\), a word \(u\), and a state \(q\), relocating the canonical one-word configuration of a single-tape phase (state \(q\), word \(u\)) into that slot, with every other tape holding its layout word at head position \(0\), yields the canonical host configuration whose layout equals the given layout with the slot’s word replaced by \(u\).
Relocating a comparator configuration at zero displacement through the comparator-pair selection yields the host’s canonical configuration whose layout holds the two whole compared words together with the unchanged candidate, every module scratch tape being blank.
For every \(j \le \left\lvert s\right\rvert \), \(j\) steps of the preparation machine from its start copy the first \(j\) bits of the candidate \(s\) onto the second tape, advance the native head to its saturated position for \(j\), and set the overflow flag exactly when \(\left\lvert w\right\rvert \lt j\); the candidate itself never changes.
For any input \(w\) and candidate \(s\), there is a time \(t\) with \(0 \lt t \le 5(\left\lvert s\right\rvert + \left\lvert w\right\rvert + 3)\) such that the preparation machine never visits its return phase before \(t\) and at time \(t\) sits in its completed frame with the candidate and native suffix prepared; positivity holds even for empty input and an empty candidate.
From the preparation machine’s candidate-rewind phase with the aligned candidate and copy heads at displacement \(j - 1\), for any \(j \le \left\lvert s\right\rvert \), exactly \(j+1\) steps restore both heads to the origin without changing either word; the native input position and the prepared suffix are fixed.
Let \(j\) be at most the length of the suffix word. From the preparation machine’s suffix-rewind phase with the suffix head at displacement \(j - 1\), exactly \(j+1\) steps yield the frame of the next phase with the suffix head at the origin; all three words are preserved.
For any input \(w\) and candidate \(s\), there is a time \(t \le 5(\left\lvert s\right\rvert + \left\lvert w\right\rvert + 3)\) after which the candidate-and-suffix preparation machine, started from its initial frame, presents the candidate \(s\) and the native suffix of \(w\) after position \(\left\lvert s\right\rvert \) on clean argument tapes, with every head and the native input position restored and the overflow flag recording whether \(\left\lvert w\right\rvert \lt \left\lvert s\right\rvert \).
For every \(j\) at most the length of the native suffix after position \(\left\lvert s\right\rvert \), \(j\) steps of the preparation machine in its suffix phase copy the first \(j\) bits of that suffix onto the third tape and advance the native head accordingly; the other tapes are unchanged.
Let the host’s transitions on embedded good states equal the relocated source transitions, with the selection map a left inverse of the tape indexing. If every state visited by the source before time \(t\) is good, then \(t\) host steps from the relocated configuration equal the relocation of the source’s \(t\)-step run.
Relocating a canonical state-word configuration of the suffix-length-call machine into the host layout — with the candidate and width words preserved — yields exactly the host’s canonical configuration on the full word layout with the length word installed.
Under the relocation hypotheses — the selection map inverts the tape indexing, no embedded state equals the host anchor, and host transitions on embedded good states equal the relocated source transitions — an exact source run from \(c\) to \(d\) in \(t\) steps, all of whose visited states are good, relocates to an exact host run between the relocated configurations whose whole trace avoids the anchor.
If the run from \(c\) reaches \(d\) in \(a\) anchor-free steps and the run from \(d\) avoids the anchor at all times strictly before \(b\), then the run from \(c\) avoids the anchor at all times strictly before \(a + b\), even when time \(a + b\) itself returns to the anchor.
The transform relocating an action of a \(k\)-tape machine to an arbitrary fixed set of slots among \(l\) host tapes, specified by a partial selection map: selected tapes perform the original writes and moves, and every unselected tape is stationary.
Given clean-call machines \(C\) (width evaluator) and \(D\) (suffix-length evaluator) with designated entry and exit states, the native controller of one width-split round: it prepares the candidate copy and native suffix, runs the two relocated calls on their disjoint banks, compares the two returned words whole-word, erases both result words while carrying the verdict, and then either emits the native split or advances the candidate and returns to the anchor; a past-the-end preparation bypasses both evaluators, the native input is never overwritten, and no time bound or width function appears in the transition table.
The host configuration embedding a \(k\)-tape configuration into \(l\) host tapes via a selection map: selected slots carry the source tapes and heads, unselected slots keep given contents and head positions, and the output is the source’s physical output.
Given a physical input \(w\), a word \(u\), a control state \(q\), and a displacement \(h\), the eraser configuration is the configuration of the administrative-word eraser in state \(q\) with \(u\) on its single tape, its head at displacement \(h\), the physical input head on the first input cell, and empty output.
A three-state, one-work-tape machine that erases one contiguous administrative word: it scans the word to its right blank, erases it right to left, and crosses the left blank once to restore the head to the origin, stalling in an absorbing return state; blanks may serve as word boundaries because the word is complete and canonical.
A configuration of the candidate-and-suffix preparation machine: the three tapes hold the candidate, its copy, and the scanned native suffix, with the candidate and copy heads aligned, a separate suffix head, a control phase with overflow flag, and a given native input position.
A three-tape machine with nine control phases that prepares the arguments of the width evaluation: it copies the candidate from its first tape onto a fresh second tape while advancing the native input head with saturation, copies the remaining native suffix onto a fresh third tape, rewinds all work heads along the preserved words and then rewinds the native input; a finite flag records only whether the candidate ran past the native right boundary.
For all positive bank sizes \(k\) and \(l\), the administrative-triple injection sends index \(0\) to the candidate position \(0\), index \(1\) to host position \(1\) (the first width-bank tape), and index \(2\) to host position \(k + 1\) (the first length-bank tape).
Given state types \(S\) and \(T\) of the width and length call machines, the controller-phase type of the width-split round has an anchor, a preparation phase, dedicated start states and running phases (carrying a state of \(S\), respectively \(T\)) for the width call and the suffix-length call, a comparison phase, two erasure phases carrying the verdict, an advance (cleanup) phase, and an emission phase.
The layout of stored words across the \(k + l + 1\) host tapes: the preserved candidate first, then the width-call bank of \(k\) tapes, then the suffix-length-call bank of \(l\) tapes; at a seam only the first tape of each bank carries a word and all other tapes are blank.
Given a machine \(M\), the right-boundary wrapper is the machine that simulates \(M\) exactly and, after the actual halt of \(M\), moves the input head to the right input boundary without altering the output or work tapes of \(M\); the initial positive move of the scan handles the two distinct blanks of an empty input correctly.
Given three families \(D\), \(V\), \(O\), each indexed by \(k\) positions, the three-bank layout is the family indexed by \(k + (k + k)\) positions that lists \(D\) first, then \(V\), then \(O\); it is used to place simulated data, visited-cell markers, and origin markers in corresponding slots.
For every function \(f\) and every word \(w\), searching the indices \(0, \dots , \left\lvert w\right\rvert \) for the first \(i\) at which the width-parametric acceptance test succeeds on the \(i\)-th iterate of the split step (the unary candidate \(1^i\)) returns exactly the width-parametric split search for \(f\) at length \(\left\lvert w\right\rvert \).
Let a body machine with an anchor state satisfy a startup contract and a round contract with common envelope \(A\, (T_E(n+1) + n + 2)\), where the round from a candidate \(s\) of length at most \(n+1\) either halts with the encoded split of the input at \(\left\lvert s\right\rvert \) when \(\left\lvert s\right\rvert + f(\left\lvert s\right\rvert )\) equals the input length, or returns to the anchor with the stepped candidate, never visiting the anchor strictly inside. Then there exist a machine \(M\) and a constant \(c\) computing, within \(c\, (n+1)(T_E(n+1) + n + 2)\) steps, the function sending \(w\) to the pair encoding of the split of \(w\) at the least \(i\) with \(i + f(i) = \left\lvert w\right\rvert \), and to the empty word when no such \(i\) exists.
Splitting \(w\) at the length of the first orbit candidate accepted by the width-parametric test and pair encoding the two halves gives the same word as splitting \(w\) at the index returned by the width-parametric split search for \(f\); when both searches fail, both sides are the empty word.
The acceptance test of the width-parametric split search: for a width function \(f\), a candidate \(s\) is accepted for the word \(w\) exactly when \(\left\lvert s\right\rvert + f(\left\lvert s\right\rvert ) = \left\lvert w\right\rvert \); no monotonicity of \(f\) is assumed.
If the input splits as \(x = p\, b\, r\) with bit \(b\), two transitions of the token encoder from its scanning state at position \(\left\lvert p\right\rvert \) emit \(b\) twice and advance the input head once; the encoder continues token scanning when \(b = 1\) and passes to the separator phase when \(b = 0\).
Let the unconsumed input \(r\) split at its leading unary token into the token \(a\) (the maximal leading run of ones with its terminating zero, if present) and the remainder \(b\). Then, after any consumed prefix and prior output \(o\), exactly \(2\left\lvert a\right\rvert + \left\lvert b\right\rvert + 3\) steps of the token encoder halt with output \(o\) followed by the self-delimiting pair encoding of \((a, b)\), the final step being the blank-reading halt.
For every input \(x\), position \(i \le \left\lvert x\right\rvert \), and prior output \(o\), two transitions of the token encoder from its separator phase at position \(i\) yield the suffix-copying state at the same position \(i\), with output \(o\) followed by the pair delimiter bits \(0,1\).
A six-state machine without work tapes that splits off its input’s leading unary token in one streaming pass: each token bit is emitted doubled, the token’s terminating \(0\) is followed by the separator bits \(0,1\), and the remainder is copied verbatim; a right blank during token scanning starts the separator directly, so an unterminated token and the empty input are both retained.
The configuration of the tracked evaluator after a completed source step: the data bank holds the source configuration’s work tapes unchanged, each visited-span marker covers the corresponding current source head, and each origin tape carries its single marker.
The configuration of the tracked evaluator between the two microsteps of a round: the source data, input position, heads, and emitted output already reflect the source action, while the visited-interval markers are still the previous ones.
Given a machine \(M\), the tracked evaluator that simulates \(M\) using two native steps per source step: an initialization step marks each origin in both marker banks, the first microstep of each round performs \(M\)’s action (moving the corresponding marker heads along), and the second stamps the newly reached head cells in the visited bank before possibly halting; the physical output is exactly \(M\)’s output.
For every control embedding, return state, output prefix \(p\), source action \(a\), and source configuration \(c\): applying the padded forwarding action of \(a\) to the forwarded configuration of \(c\) equals the forwarded configuration of \(a\) applied to \(c\), appending the optional emitted bit after the accumulated prefix, including on a halting action.
Let a source machine and a host machine be given, together with a control embedding and a return state, such that on every embedded state the host transition is the padded forwarding of the source transition. If the source run from a configuration \(c\) stays live before time \(t\), then the host run for \(t\) steps from the forwarded configuration of \(c\) with prefix \(p\) equals the forwarded configuration of the source run at time \(t\) with the same prefix; the source may halt on the final action, whose emission is forwarded.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet, \(c\) a configuration, \(p\) a word, and \(t\) a time. Running \(M\) for \(t\) steps from \(c\) with \(p\) prepended to its output equals the run of \(M\) from \(c\) at time \(t\) with \(p\) prepended to its output, including halted tails.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet, \(c\) a configuration, and \(p\) a word. One step of \(M\) from \(c\) with \(p\) prepended to its output equals the one-step successor of \(c\) with \(p\) prepended to its output: transitions never inspect the append-only output.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet, let \(c_0, \dots , c_{N-1}\) be configurations on a common input with empty outputs, let \(u_0, \dots , u_{N-1}\) be chunk words, and suppose \(N \gt 0\) and that for each \(i \lt N\), within \(B\) steps, the machine run from \(c_i\) reaches \(c_{i+1}\) with output exactly \(u_i\) when \(i + 1 \lt N\), and halts with output exactly \(u_i\) when \(i + 1 = N\). Then for every accumulated prefix \(p\), the machine run from \(c_0\) with output \(p\) halts within \(N \cdot B\) steps with output \(p\) followed by the concatenation \(u_0 u_1 \cdots u_{N-1}\) in order.
For an output prefix \(p\), a startup marker, a body configuration \(c\), a release bit, an optional flag bit, a counter word, and a fuel configuration, the forwarding-host configuration whose body, counter, and fuel data are those of the padded stopped-body configuration under the forwarding discipline, whose physical output is \(p\) followed by the output of \(c\), and whose now-unused payload tape is blank with head at the origin.
For every startup marker, body configuration \(c\) with empty output, release bit, flag bit, counter word, and fuel configuration, the forwarding call on \(c\) with empty prefix equals the corresponding capturing call on \(c\).
For every output prefix \(p\), startup marker, body configuration \(c\), release bit, flag bit, counter word, and fuel configuration: the forwarding call equals the corresponding capturing call built on \(c\) with its output removed, with physical output replaced by \(p\) followed by the output of \(c\). The identity is between configurations only.
For a control embedding, a return state, an accumulated output prefix \(p\), and a source configuration \(c\), the host configuration that embeds \(c\) with one extra blank inactive tape, carries \(p\) followed by the output of \(c\) as physical output, and embeds the control of \(c\), using the return state when \(c\) has halted.
Given a loop body with anchor, a fuel machine, and a mode bit, the variant of the loop controller whose body calls route every body emission directly to the physical output (instead of capturing it), while the fuel capture, counter countdown, and all administrative transitions are those of the original controller.
Let \(c\) be a body configuration, \(p\) an output prefix, and \(t\) a time at which the body run from \(c\) sits in the anchor state, with no earlier anchor visit except possibly a released step at time zero and with the release bit clear whenever \(t = 0\). Then the forwarding host, run for \(t + 1\) steps from the forwarding call on \(c\) with prefix \(p\), reaches the forwarding stopped call on the body run at time \(t\) with control removed, release clear, and stop flag \(\mathrm{false}\), the counter word and fuel residue unchanged.
Let \(c\) be a configuration of the extended stopped-body source on the forwarding host’s layout, \(p\) an output prefix, and \(t\) a time before which that source does not halt. Then the forwarding host run for \(t\) steps from the forwarded embedding of \(c\) with prefix \(p\) equals the forwarded embedding of the source run at time \(t\) with the same prefix.
Let \(w\) be a counter word and \(j\) the length of its maximal \(\mathrm{false}\) prefix. Then \(2j + 2 \le 2|w| + 2\), and running the forwarding host for exactly \(2j + 2\) steps from its borrow phase on a frame whose counter track is the buffer tape of \(w\) with head at zero ends with the counter track holding the decremented word, head at zero, in the released next-anchor call state on success and in the exhaustion phase on underflow.
Let \(w\) be a counter word, \(b\) a success bit, and \(j \le |w|\). Running the forwarding host for exactly \(j + 1\) steps from the rewind phase selected by \(b\) on a frame whose counter track is the buffer tape of \(w\) with head at cell \(j - 1\) returns the counter head to zero and dispatches: success releases the next anchor call, underflow enters the exhaustion phase without yet emitting.
Let \(p\) and \(w\) be Boolean words and \(j\) the length of the maximal \(\mathrm{false}\) prefix of \(w\). Running the forwarding host for exactly \(j + 1\) steps from its borrow phase on a frame whose counter track holds \(p\) followed by \(w\) with head at \(|p|\) reaches the success- or underflow-rewind phase according to the decrement’s flag, with the counter track holding \(p\) followed by the decremented word and head at \(|p| + j - 1\).
Let \(p\) and \(r\) be Boolean words. One step of the forwarding host from its borrow phase on a frame whose counter track is the buffer tape of \(p\) followed by \(r\) with counter head at \(|p|\) behaves as the standalone decrement machine: an empty \(r\) enters the underflow rewind, a leading \(\mathrm{true}\) writes \(\mathrm{false}\) and enters the success rewind, and a leading \(\mathrm{false}\) writes \(\mathrm{true}\) and continues the borrow. Only the counter track changes.
Let \(c\) be a configuration of the relocated fuel source on the forwarding host’s tape layout and \(t\) a time before which that source does not halt. Then the forwarding host run for \(t\) steps from the capture embedding of \(c\) equals the capture embedding of the fuel source run at time \(t\).
Let \(p\) be the already-copied prefix and \(r\) the remaining fuel suffix. Starting the forwarding host in its counter-copy phase on a controller frame whose counter track is the buffer tape of \(p\) and whose payload track is the partially cleared copying tape, with both heads at cell \(|p|\), after exactly \(|r| + 1\) steps the host is in the synchronized-rewind phase with the counter track holding the full fuel word, the payload track blank, and both heads one cell left of the word’s right end.
Let \(w\) be the copied counter word and \(j \le |w|\). Starting the forwarding host in its synchronized-rewind phase on a controller frame whose counter track is the buffer tape of \(w\) and whose payload track is blank, with both heads at cell \(j - 1\), after exactly \(j + 1\) steps the host is in the input-rewind phase with both heads at cell zero and every other frame component unchanged.
Let \(w\) be a stored word and \(j \le |w|\). Starting the forwarding host in its fuel- rewind phase on a controller frame whose payload track is the buffer tape of \(w\) with payload head at cell \(j - 1\), after exactly \(j + 1\) steps the host is in the counter- copy phase with the payload head at cell zero and every other frame component unchanged.
Let \(w\) be the captured fuel word. Starting the forwarding host in controller phase zero on a frame whose counter track is blank and whose payload track is the buffer tape of \(w\) with payload head at cell \(|w|\), after exactly \(3|w| + 4\) steps the host is in the input-rewind phase with the counter track holding \(w\), the payload track blank, both heads at cell zero, and all inactive data unchanged.
Let \(c\) be a configuration of the forwarding host in controller phase four. Then there is a time \(t\) at most the input-head position of \(c\) plus two after which the host sits in the startup body-call state with input head at position one and every other component of \(c\) unchanged.
Let the fuel machine compute, in time \(T\), the function sending each input \(x\) to the binary digit word of \(R(|x|)\). Then for every input \(x\) there are a halted fuel configuration \(c\) with output exactly that digit word and a time \(t \le 5\, T(|x|) + 7\) such that the forwarding host, run for \(t\) steps from its initial configuration on \(x\), reaches the prepared startup configuration built from \(c\).
Let \(c\) be a halted body configuration with empty output, \(w\) a counter word, and \(f\) a fuel configuration. Then there is a time \(t \le 2|w| + 4\) such that, from the stopped call on \(c\) with stop flag \(\mathrm{false}\): if the fixed-width decrement of \(w\) succeeds, the forwarding host at time \(t\) is at the released next-anchor call with the decremented counter word; otherwise the host at time \(t\) has halted with output empty in find mode and the single bit \(\mathrm{false}\) in decision mode.
Let \(c\) be a body configuration, \(w\) a counter word, and \(f\) a fuel configuration. One step of the forwarding host from the startup call on \(c\) with control removed, release clear, and flag \(\mathrm{false}\) yields the active call on \(c\) with control the anchor, release set, and no flag, changing no body, counter, or fuel data.
Let \(s\) be a round word, \(w\) a counter word, \(f\) a fuel configuration, and \(t \gt 0\) a time such that the body, started at the clean anchor seam carrying \(s\), does not revisit the anchor strictly between times zero and \(t\) and at time \(t\) sits at the clean seam carrying the stepped word with output a chunk word \(u\). Then there is a time \(v \le t + 2|w| + 5\) such that from the released call on the seam of \(s\): when the fixed-width decrement of \(w\) succeeds, the forwarding host at time \(v\) is at the released call on the next seam with the decremented counter and physical output exactly \(u\); on underflow the host at time \(v\) has halted with output exactly \(u\) and no verdict bit.
Let \(s\) be a word, \(f\) a fuel configuration, and \(t\) a time such that the body, run from its initial configuration, first visits the anchor at time \(t\), arriving at the clean seam that carries \(s\) on work tape zero with all other work tapes blank. Then the forwarding host, run for \(t + 2\) steps from the prepared startup configuration, reaches the released call on that seam with counter word the output of \(f\) and fuel residue \(f\), with nothing emitted and no fuel debited.
Let \(M\) and \(U\) be multi-tape Turing machines over the Boolean alphabet, let \(f\) and \(g\) be functions on Boolean words, and let \(T_1, T_2 : \mathbb {N} \to \mathbb {N}\) be arbitrary. Suppose \(M\) computes \(f\) in time \(T_1\), meaning that on every input \(x\) it halts within \(T_1(\left\lvert x\right\rvert )\) steps with output \(f(x)\), and suppose that for every Boolean word \(x\) the machine \(U\) on input \(f(x)\) halts within \(T_2(\left\lvert x\right\rvert )\) steps with output \(g(x)\), the budget being measured at the length of the original word \(x\). Then there is a machine \(N\) over the Boolean alphabet that computes \(g\) in time \(n \mapsto 2\, T_1(n) + T_2(n) + 2\); no monotonicity of \(T_2\) is required, and \(U\) need not halt on words outside the image of \(f\).
For any two multi-tape Turing machines \(M_1\) and \(M_2\) over the Boolean alphabet — with no totality or time hypotheses — there exists a machine \(M\) such that for all Boolean words \(x\) and \(w\): \(M\) halts on input \(x\) with output \(w\) (within some number of steps) if and only if there is a Boolean word \(y\) such that \(M_1\) halts on input \(x\) with output \(y\) and \(M_2\) halts on input \(y\) with output \(w\). In particular, if \(M_1\) diverges on \(x\), or halts but \(M_2\) diverges on its output, then \(M\) diverges on \(x\).
Let \(D\), \(M_1\), \(M_2\) be multi-tape Turing machines over the Boolean alphabet and let \(p\) be a Boolean-valued predicate on Boolean words such that \(D\) computes the function \(x \mapsto [p(x)]\): on every input \(x\) it eventually halts with the single symbol \(p(x)\) as its output. Then there exists a machine \(M\) such that for all Boolean words \(x\) and \(w\): \(M\) halts on input \(x\) with output \(w\) (within some number of steps) if and only if the branch machine selected by \(p(x)\) — namely \(M_1\) when \(p(x)\) is true and \(M_2\) when \(p(x)\) is false — halts on that same input \(x\) with output \(w\) (within some number of steps).
Let \(M\) be a bundled finite machine computing a function \(f\) on Boolean words in time \(T\). Then there are a machine \(C\) with a positive number of work tapes, control states called entry and exit, and a constant \(c\) such that for every native input \(x\) and argument \(a\): starting \(C\) at the canonical seam in the entry state with \(a\) as the sole tape-resident word, there is a positive time \(t \le c\, (T(|a|) + |a| + |f(a)| + 1)\), before which the exit state is never visited, at which \(C\) is exactly at the entry seam with exit control and physical output \(f(a)\): the argument preserved on its tape, all scratch restored, and the computed word emitted.
Let a loop body with anchor be given, a fuel machine computing the binary digits of \(R(|x|)\) in time \(T\), an invariant that holds at the initial word and is preserved by the step function, a startup reaching the initial seam within \(T(|x|)\) without visiting the anchor earlier, and a per-round contract: on every admissible word \(s\), in positive time at most \(T(|x|)\) with no early anchor revisit, the body moves from the clean seam of \(s\) to the clean seam of the stepped word while appending exactly the chunk word assigned to \((x, s)\) to its output. Then there are a finite machine \(E\) and a constant \(c\) such that \(E\) computes, within \(c (T(n) + 1)(R(n) + 2)\) steps on inputs of length \(n\), the concatenation of the chunks of the first \(R(|x|) + 1\) orbit points of the initial word, in order.
Let \(M\) be a bundled finite machine computing a function \(f\) on Boolean words in time \(T\). Then there are a machine \(C\) with a positive number of work tapes, control states called entry and exit, and a constant \(c\) such that for every native input \(x\) and argument \(a\): starting \(C\) at the canonical seam in the entry state with \(a\) as the sole tape-resident word (all other tapes blank, input head at one, empty output), there is a positive time \(t \le c\, (T(|a|) + |a| + |f(a)| + 1)\), before which the exit state is never visited, at which \(C\) is exactly at the canonical seam in the exit state carrying \(f(a)\) as the sole tape-resident word: the result installed, all scratch restored, nothing emitted, and the native input untouched.
Let a loop body with anchor, a fuel machine computing the binary digits of \(R(|x|)\) in time \(T\), an invariant preserved from the initial word by the step function, a bounded anchor-free startup, and a per-round accept-or-advance contract with verdict output the single bit \(\mathrm{true}\) be given. Then there are a finite machine \(E\) and a constant \(c\) such that on every input \(x\), \(E\) has a family of round configurations: a startup of at most \(c (T(|x|) + 1)\) steps reaching round zero, empty output at rounds \(0, \dots , R(|x|)\), a halted terminal at index \(R(|x|) + 1\) with output the single bit \(\mathrm{false}\), and per-round segments of at most \(c (T(|x|) + 1)\) steps that halt with \(\mathrm{true}\) on acceptance and advance to the next round configuration otherwise.
Let a loop body with anchor be given, a fuel machine computing the binary digits of \(R(|x|)\) in time \(T\), an invariant that holds at the initial word and is preserved by the step function, a bounded anchor-free startup, and a per-round contract in which an accepting round halts with a declared payload word and a rejecting round advances to the seam of the stepped word. Then there are a finite machine \(E\) and a constant \(c\) such that \(E\) computes, within \(c (T(n) + 1)(R(n) + 2)\) steps on inputs of length \(n\), the payload of the first accepted orbit point among the first \(R(|x|) + 1\) iterates of the initial word, and the empty word when no orbit point is accepted.
Let a loop body with anchor be given, a fuel machine computing the binary digits of \(R(|x|)\) in time \(T\), an invariant that holds at the initial word and is preserved by the step function, a startup reaching the initial seam within \(T(|x|)\) without visiting the anchor earlier, and a per-round contract: on every admissible word, in positive time at most \(T(|x|)\) with no early anchor revisit, the body either halts with output the single bit \(\mathrm{true}\) or advances to the seam of the stepped word. Then there are a finite machine \(E\) and a constant \(c\) such that \(E\) computes, within \(c (T(n) + 1)(R(n) + 2)\) steps on inputs of length \(n\), the one-bit answer to whether some orbit point among the first \(R(|x|) + 1\) iterates of the initial word is accepted.
If the input splits as \(x = p\, b\, c\, r\) with bits \(b, c\), then two extractor steps from the block-start state at position \(\left\lvert p\right\rvert \) with buffer \(a\) append \(b\) to the buffer and continue when \(b = c\); halt silently when \((b,c) = (1,0)\); and enter the rewind phase with the work head one cell before the end of the buffer when \((b,c) = (0,1)\). In all three cases nothing is emitted.
For a split \(x = p s\) of the input and a validated buffer \(a\), there is a time \(t \le 2\left\lvert a\right\rvert + \left\lvert s\right\rvert + 3\) such that after \(t\) steps from the start of the rewind phase at position \(\left\lvert p\right\rvert \) the extractor has halted, and its output is \(a\) (when the first flag is set) followed by \(s\) (when the second flag is set).
If the input splits as \(x = p\, b\, r\) with bit \(b\), one step of the pair extractor from its block-start state at position \(\left\lvert p\right\rvert \) records \(b\) in the finite control and advances the input head, leaving the buffer and the output unchanged.
From the extractor’s replay state with buffer \(a\) and work head at displacement \(0\), for any \(j \le \left\lvert a\right\rvert \), exactly \(j\) steps move the work head to displacement \(j\) and emit the first \(j\) bits of \(a\) when the first-component flag is set, and nothing otherwise.
From the extractor’s replay state with buffer \(a\) and work head at displacement \(0\), exactly \(\left\lvert a\right\rvert +1\) steps reach the suffix state with the work head at displacement \(\left\lvert a\right\rvert \), and the output is \(a\) when the first-component flag is set and empty otherwise.
From the extractor’s rewind state with buffer \(a\) and work head at displacement \(j-1\), for any \(j \le \left\lvert a\right\rvert \), exactly \(j+1\) steps return the work head to displacement \(0\) and enter the replay state, with the input position, the buffer, and the (empty) output unchanged.
For any split \(x = p s\) of the input and any buffered prefix \(a\), there is a time \(t \le 3\left\lvert s\right\rvert + 2\left\lvert a\right\rvert + 5\) such that after \(t\) steps from the block-start state at position \(\left\lvert p\right\rvert \) the extractor has halted; if \(s\) decodes under the self-delimiting pair code as \((b,c)\), the output is \(a b\) (when the first flag is set) followed by \(c\) (when the second flag is set), and if \(s\) is not a valid pair code the output is empty.
With the second-component flag set and a split \(x = p s\) of the input, from the extractor’s suffix state at position \(\left\lvert p\right\rvert \) with buffer \(a\) and prior output \(o\), exactly \(\left\lvert s\right\rvert +1\) steps reach a halted configuration with output \(o\) followed by \(s\); the buffer and its head are untouched.
Given an input \(x\), an optional control state, an input position \(i \le \left\lvert x\right\rvert \), a buffered word \(a\), a work-head displacement \(z\), and an output word, the extractor configuration is the configuration of the pair extractor with that control state, with its input head past the first \(i\) cells, with \(a\) stored on its single work tape and the work head at \(z\), and with the given output.
In an extractor configuration at input position \(i \le \left\lvert x\right\rvert \), the symbol under the input head is the \(i\)-th bit of \(x\) when \(i \lt \left\lvert x\right\rvert \) and the blank otherwise, regardless of the buffer contents.
Given, for each of \(k\) work tapes, the current side of the fold and the physical cell under the head, this translates an action of a \(k\)-work-tape machine over a base alphabet \(\Gamma \) with state set \(S\) into an action over the folded alphabet whose states pair a state of \(S\) with a side for each tape (or are absent, meaning halt): writes become side-relative folded writes, moves are translated across the fold using the origin flags stored in the scanned cells, the emitted output symbol is embedded, and the successor state is paired with the updated sides.
Let \(c\) be a configuration of a \(k\)-work-tape machine over the folded alphabet, on an arbitrary folded-alphabet input, satisfying the safety invariant (control not in the fresh initial state, nonnegative heads, correct origin flags at all nonnegative cells). Then for every assignment of sides to the work tapes and every base-machine action \(a\), applying the translation of \(a\) — relative to those sides and the cells under the heads of \(c\) — again yields a configuration satisfying the safety invariant.
The folded image of a configuration \(c\) of a \(k\)-work-tape machine over a base alphabet \(\Gamma \) with state set \(S\) on an input \(x\): a configuration over the folded alphabet on the symbol-wise embedded input, whose control state is the state of \(c\) paired with the side of each work head, whose work tapes are the folded tapes of those of \(c\), whose work head positions are the folded coordinates \(\varphi \) of those of \(c\), and whose input head position and output are transported along the embedding.
Let \(\Gamma \) be an alphabet with decidable equality, \(a\) an action of a \(k\)-work-tape machine over \(\Gamma \) with state set \(S\), and \(c\) a configuration of such a machine on an input \(x\). Applying to the folded configuration of \(c\) the translated action — formed from each work head’s current side and the physical cells under the folded heads — yields the folded configuration of the result of applying \(a\) to \(c\).
Let \(M\) be a multi-tape machine over a finite alphabet \(\Gamma \) with decidable equality and let \(x\) be an input over \(\Gamma \). One step of the folding simulator from its initial configuration on the symbol-wise embedded input \(x\) yields exactly the folded configuration of the initial configuration of \(M\) on \(x\).
For every configuration \(c\) of a \(k\)-work-tape machine over a base alphabet \(\Gamma \) on an input \(x\), the possibly absent input symbol scanned in the folded configuration of \(c\) equals the image under the alphabet embedding of the input symbol scanned in \(c\) (with absence mapped to absence).
In the folded image of a configuration \(c\) of a \(k\)-work-tape machine over a base alphabet, for each work tape the origin flag obtained by unpacking the physical cell under that tape’s head equals the truth value of the statement that the folded head coordinate of that tape is \(0\).
Let \(M\) be a multi-tape machine over a finite alphabet \(\Gamma \) with decidable equality, \(x\) an input over \(\Gamma \), and \(t \in \mathbb {N}\). Running the folding simulator for \(t+1\) steps from its initial configuration on the symbol-wise embedded input equals the folded configuration of running \(M\) for \(t\) steps from its initial configuration on \(x\).
Let \(M\) be a multi-tape machine over a finite alphabet \(\Gamma \) with decidable equality and let \(c\) be a configuration of \(M\) on an input \(x\) over \(\Gamma \). One step of the folding simulator of \(M\) on the folded configuration of \(c\) equals the folded configuration of one step of \(M\) on \(c\).
The single initialization action of the folding simulator, parametrized by a target base state \(q\): without moving any head and without emitting output, it simultaneously writes the origin marker — the folded triple with flag set and both payloads blank — at the current cell of every work tape, and passes control to the state \(q\) with every tape recorded as being on the nonnegative side.
Given the current side of the fold, an origin flag, and a simulated move direction \(d \in \{ -1, 0, +1\} \), this operation returns the physical move to perform together with the new side: moves translate directly on the nonnegative side and are reversed on the negative side, while a crossing between simulated cells \(0\) and \(-1\) (detected by the origin flag) flips the side without issuing a physical move.
Let \(z\) be an integer and \(d \in \{ -1,0,+1\} \) a move direction; write \(\varphi (z)\) for the folded coordinate (\(z\) when \(z \ge 0\), else \(-z-1\)) and \(\sigma (z)\) for the side of \(z\) (whether \(z \ge 0\)). Then the translated move computed from \(\sigma (z)\), the flag “\(\varphi (z) = 0\)”, and \(d\) is correct in both components: adding its physical displacement to \(\varphi (z)\) gives \(\varphi (z+d)\), and its returned side equals \(\sigma (z+d)\).
The canonical packing of a folded-alphabet triple into an optional cell content: the untagged triple with both payloads blank is sent to the physical blank, and every other triple is kept as an actual symbol. This ensures that a physical cell never stores a symbol carrying no information.
The coordinate-folding map \(\varphi \colon \mathbb {Z}\to \mathbb {Z}\) defined by \(\varphi (z) = z\) when \(z \ge 0\) and \(\varphi (z) = -z-1\) otherwise. It maps each negative cell \(z\) onto the nonnegative cell \(-z-1\); in particular \(\varphi (0) = \varphi (-1) = 0\), so \(\varphi \) is not the absolute value.
Reading a physical cell over the folded alphabet relative to a side: after unpacking the cell (a physical blank unpacks to the untagged triple of blanks), return the first payload when the side is the nonnegative one and the second payload otherwise.
For every tape content \(t \colon \mathbb {Z}\to \Gamma \cup \{ \square \} \) over a base alphabet \(\Gamma \) and every integer \(z\), reading the folded tape of \(t\) at the physical position \(\varphi (z)\) on the side of \(z\) returns exactly the original content \(t(z)\), where \(\varphi (z) = z\) for \(z \ge 0\) and \(\varphi (z) = -z-1\) otherwise.
The safety invariant on a configuration \(c\) of a \(k\)-work-tape machine over the folded alphabet whose states are either a fresh initial state or a base state paired with per-tape sides, on an arbitrary folded-alphabet input: the control state of \(c\) is not the fresh initial state (it may be halted), every work head position is nonnegative, and on every work tape the origin flag obtained by unpacking each nonnegative cell \(p\) is set exactly when \(p = 0\).
Let \(M\) be a multi-tape machine over a finite alphabet with decidable equality. For every input over the folded alphabet — including inputs containing symbols outside the range of the alphabet embedding — the configuration reached after one step of the folding simulator from its initial configuration satisfies the safety invariant.
Let \(M\) be a multi-tape machine over a finite alphabet with decidable equality and let \(c\) be a configuration of the folding simulator of \(M\), on an arbitrary folded-alphabet input, satisfying the safety invariant. Then the configuration obtained after one step of the simulator also satisfies the safety invariant.
The action taken from the simulator’s fresh initial state when scanning a possibly absent folded input symbol: if the scanned symbol is absent or decodes to a base symbol, perform the initialization action toward a given base state \(q\); if it lies outside the range of the alphabet embedding, perform the same origin-marking writes but halt instead of entering \(q\). Thus no simulated transition precedes the validity check on the first input symbol.
For every base state \(q\) and every possibly absent base symbol, the guarded start action evaluated at the image of that symbol under the alphabet embedding equals the plain initialization action toward \(q\): embedded symbols never trigger the halting guard.
The enlarged tape alphabet used to fold a two-way infinite tape over a base alphabet \(\Gamma \): its symbols are triples \((b, u, v)\) where \(b\) is a Boolean origin flag and \(u\) and \(v\) are two independent payloads, each either a symbol of \(\Gamma \) or blank. The intent is that a nonnegative physical cell \(p\) carries the content \(u\) of the simulated cell \(p\) and the content \(v\) of the simulated cell \(-p-1\).
Given a two-way infinite tape content \(t \colon \mathbb {Z}\to \Gamma \cup \{ \square \} \) over a base alphabet \(\Gamma \) (with \(\square \) the blank), the folded tape assigns to each nonnegative physical cell \(p\) the canonical packing of the triple consisting of the flag “\(p = 0\)”, the content \(t(p)\), and the content \(t(-p-1)\); every negative physical cell is blank.
Let \(\Gamma \) be an alphabet with decidable equality, \(t \colon \mathbb {Z}\to \Gamma \cup \{ \square \} \) a tape content, \(z\) an integer, and \(a\) a possibly blank symbol. Updating the folded tape of \(t\) at the physical cell \(\varphi (z)\) by the side-relative write of \(a\) (on the side of \(z\)) yields exactly the folded tape of the pointwise update of \(t\) at \(z\) by \(a\); here \(\varphi (z) = z\) for \(z \ge 0\) and \(\varphi (z) = -z-1\) otherwise.
For a multi-tape machine \(M\) over a finite alphabet \(\Gamma \) with decidable equality, the simulating machine over the folded alphabet: it has the same number of work tapes as \(M\), and its states are either a fresh initial state or a state of \(M\) paired with a side for each work tape. From the fresh state it performs the guarded initialization; thereafter it decodes the scanned input symbol (executing the stationary halt if the symbol lies outside the embedding’s range), reads the active payload of each work tape, and executes the translation of the corresponding transition of \(M\).
For every multi-tape machine \(M\) over a finite alphabet with decidable equality, the folding simulator of \(M\) has nonnegative heads: for every input over the folded alphabet, every number of steps, and every work tape, the position of that tape’s head in the corresponding initialized run is nonnegative.
Writing a possibly blank base symbol \(a\) into a physical cell over the folded alphabet, relative to a side: unpack the cell, replace the first payload by \(a\) when the side is the nonnegative one and the second payload otherwise, keep the origin flag and the untouched payload, and repack canonically (so a fully blank untagged result becomes a physical blank).
For every side \(s\), every physical cell content \(w\) over the folded alphabet, and every possibly blank base symbol \(a\), the origin flag obtained by unpacking the result of the side-relative write of \(a\) into \(w\) equals the origin flag obtained by unpacking \(w\).
The machine over the Boolean alphabet with no work tapes and a single state that copies its input to its output: on scanning an input symbol it appends that symbol to the output and moves the input head one cell right, and on scanning a boundary blank it halts.
Let \(x\) be a Boolean word and run the one-state copy machine from its initial configuration on input \(x\). For every \(t \le \left\lvert x\right\rvert \), after \(t\) steps the machine is still live in its unique state, its input head sits at position \(t + 1\) (in the convention where position \(0\) is the left boundary blank and position \(j\) with \(1 \le j \le \left\lvert x\right\rvert \) scans the \(j\)-th input symbol), and its output is exactly the first \(t\) symbols of \(x\).
Fix Boolean words \(w_0\), \(u\), \(v\), an input word \(x\), and a Boolean \(b\), and write \(z\) for the word \(u\) if \(b\) is true and \(v\) if \(b\) is false. Let \(c\) be any configuration of the fixed-word comparator machine for \(w_0\), \(u\), \(v\) on input \(x\) whose state is the initial state of the emission chain for \(z\) and whose output is still empty. Then after exactly \(\left\lvert z\right\rvert + 1\) further steps the machine has halted with output exactly \(z\), independently of the position of the input head in \(c\).
Fix Boolean words \(w_0\), \(u\), \(v\) and an input word \(x\), and let \(i\) and \(r\) be natural numbers with \(\left\lvert w_0\right\rvert = i + r\), \(i \le \left\lvert x\right\rvert \), and the first \(i\) symbols of \(x\) equal to the first \(i\) symbols of \(w_0\). From any configuration of the fixed-word comparator machine for \(w_0\), \(u\), \(v\) on input \(x\) that is in comparison state \(i\) with its input head at position \(i + 1\) (position \(0\) being the left boundary blank) and with empty output, there is a \(t \le r + \max (\left\lvert u\right\rvert , \left\lvert v\right\rvert ) + 2\) such that after \(t\) steps the machine has halted, with output \(u\) if \(x = w_0\) and output \(v\) otherwise.
For fixed Boolean words \(w_0\), \(u\), \(v\), the machine over the Boolean alphabet with no work tapes whose states are \(\left\lvert w_0\right\rvert + 1\) comparison states together with an emission chain for \(u\) and an emission chain for \(v\). Comparison state \(i \lt \left\lvert w_0\right\rvert \) moves the input head right and passes to comparison state \(i + 1\) when the scanned symbol is the \(i\)-th symbol of \(w_0\), and enters the emission chain for \(v\) otherwise; comparison state \(\left\lvert w_0\right\rvert \) enters the chain for \(u\) on a boundary blank and the chain for \(v\) on any remaining symbol. Each emission chain writes its word one symbol per step and halts.
Every word \(x \in \{ 0,1\} ^*\), read least-significant bit first, satisfies exactly one of the following: \(x\) consists entirely of ones and its fixed-width increment overflows (is undefined); or \(x = 1^j 0 r\) for some \(j \ge 0\) and tail \(r\), and the fixed-width increment of \(x\) is \(0^j 1 r\).
On every input \(x\), including the empty word, the increment machine halts within \(3(\left\lvert x\right\rvert +1)\) steps; its output is the fixed-width binary increment of \(x\) (least-significant bit first) when no overflow occurs, and is empty when \(x\) consists entirely of ones.
The fixed-width increment machine is a four-state machine without work tapes computing the fixed-width binary increment of its input, read least-significant bit first: it scans the leading ones silently, and on reaching a \(0\) it rewinds and re-scans, emitting a \(0\) for each carried \(1\), a \(1\) in place of the first \(0\), and the remaining bits unchanged; an all-ones input halts with empty output.
Let \(\mathit{visit}_i \colon V \times C \to V \times C\) be a family of local rules indexed by a type \(I\) with decidable equality, let \(\mathit{cell} \colon I \to C\) assign a cell to each index, and let \(\mathit{is}\) be a list of indices with no duplicates. Transducing, with the table-valued indexed transducer started at table \(s\), the block that lists each \(i \in \mathit{is}\) tagged with \(\mathit{cell}(i)\), yields: the table that at every index \(i\) occurring in \(\mathit{is}\) holds the new value of \(\mathit{visit}_i\) applied to the original entry \(s(i)\) and \(\mathit{cell}(i)\), and elsewhere agrees with \(s\); together with the block in which each tagged cell is replaced by its rewritten value, likewise computed from the original table. Because no index repeats, each local rule sees the untouched table entry.
Let \(k \in \mathbb {N}\), let \(\mathit{visit}_i \colon V \times C \to V \times C\) be a family of local rules indexed by \(i \in \{ 0, \dots , k-1\} \), let \(\mathit{cell}\) assign a cell to each index, and let \(s\) be an initial table. Transducing, with the table-valued indexed transducer started at \(s\), the block listing every index \(i\) exactly once in increasing order tagged with \(\mathit{cell}(i)\), yields the table \(i \mapsto \) (new value of \(\mathit{visit}_i\) on \(s(i)\) and \(\mathit{cell}(i)\)) together with the block of rewritten tagged cells in the same order.
Let \(k \in \mathbb {N}\), let \(\mathit{visit}_i \colon V \times C \to V \times C\) be a family of local rules indexed by \(i \in \{ 0, \dots , k-1\} \), let \(\mathit{cell}\) assign a cell to each index, and let \(s\) be an initial table. Transducing, with the table-valued indexed transducer started at \(s\), the block listing every index exactly once in decreasing order tagged with its cell, yields the table \(i \mapsto \) (new value of \(\mathit{visit}_i\) on \(s(i)\) and \(\mathit{cell}(i)\)) together with the correspondingly reversed block of rewritten tagged cells.
Given a family of local rules \(\mathit{visit}_i \colon V \times C \to V \times C\) indexed by a type \(I\) with decidable equality, this defines a transducer whose control state is a table \(s \colon I \to V\) and whose cells are index-tagged pairs \((i, c)\): on such a pair it applies \(\mathit{visit}_i\) to the table entry \(s(i)\) and the cell \(c\), updates the table at \(i\) to the new value, and outputs \(i\) tagged with the rewritten cell. Only the entry named by the visited cell changes.
Let \(c\) be a configuration of a multi-tape Turing machine over the Boolean alphabet on input \(x\), and let \(i \le \left\lvert x\right\rvert \) be a natural number such that the input head of \(c\) sits at position \(i + 1\), in the convention where position \(0\) is the left boundary blank and position \(j\) with \(1 \le j \le \left\lvert x\right\rvert \) scans the \(j\)-th input symbol. Then the symbol scanned on the input tape is the optional \(i\)-th entry of \(x\) (zero-indexed): the symbol \(x_i\) when \(i \lt \left\lvert x\right\rvert \), and the right boundary blank when \(i = \left\lvert x\right\rvert \).
Given natural numbers \(k\) and \(l\), a state renaming \(f : S \to S'\), and a one-step action \(a\) of a machine with \(k\) work tapes over the Boolean alphabet, the action of a machine with \(k + l\) work tapes that performs \(a\) on the first \(k\) work tapes, writes nothing and moves no head on the last \(l\) tapes, keeps the input-head movement and the emitted output of \(a\), and renames the successor state along \(f\), preserving halting.
Given a state renaming \(f : S \to S'\), a configuration \(c\) of a machine with \(k\) work tapes over the Boolean alphabet on input \(x\), and arbitrary contents and head positions for \(l\) further tapes, the configuration of a machine with \(k + l\) work tapes on the same input whose state is that of \(c\) renamed along \(f\) (halted exactly when \(c\) is halted), whose input head and output agree with those of \(c\), whose first \(k\) work tapes and heads are those of \(c\), and whose last \(l\) work tapes carry the given inactive contents and head positions.
Let \(f : S \to S'\) be a state renaming, \(a\) a one-step action of a machine with \(k\) work tapes over the Boolean alphabet, \(c\) a configuration of such a machine on input \(x\), and fix contents and head positions for \(l\) inactive extra tapes. Applying the left-block extension of \(a\) along \(f\) to the left-block embedding of \(c\) with the given inactive data yields the left-block embedding, with the same inactive data, of the result of applying \(a\) to \(c\).
Let \(M\) be a machine with \(k\) work tapes and \(M'\) a machine with \(k + l\) work tapes, both over the Boolean alphabet, and let \(f\) be a map from the states of \(M\) to the states of \(M'\) such that on every renamed state, for all scanned symbols, the transition of \(M'\) is the left-block extension along \(f\) of the corresponding transition of \(M\) read off the first \(k\) tapes. Then for every configuration \(c\) of \(M\) on input \(x\), all inactive contents and head positions for the last \(l\) tapes, and every number of steps \(t\), running \(M'\) for \(t\) steps from the left-block embedding of \(c\) equals the left-block embedding of the run of \(M\) for \(t\) steps from \(c\).
Let \(M\) be a machine with \(k\) work tapes and \(M'\) a machine with \(k + l\) work tapes, both over the Boolean alphabet, and let \(f\) be a map from the states of \(M\) to the states of \(M'\) such that on every renamed state \(f(q)\), for every scanned input symbol and every tuple of scanned work symbols, the transition of \(M'\) is the left-block extension along \(f\) of the transition of \(M\) from \(q\) on that input symbol and the symbols scanned on the first \(k\) tapes. Then for every configuration \(c\) of \(M\) on input \(x\) and all inactive contents and head positions for the last \(l\) tapes, one step of \(M'\) from the left-block embedding of \(c\) equals the left-block embedding of one step of \(M\) from \(c\); in particular a halted embedded configuration stays halted.
For a machine \(M\), the administrative action of the length checker built over \(M\), given an input direction, a countdown direction, an optional bit, and a target state, is the action that moves the input head and the final (countdown) work head by those directions, emits the optional bit, passes to the target state, and writes no tape cell.
Applying an administrative action with compatible input-head and countdown-head moves to a length-checker configuration yields the checker configuration with the correspondingly updated input position and countdown counter, with output equal to the action’s optional emission and all work tapes unchanged.
Given a machine \(M\), an input \(x\), a configuration \(c\) of \(M\) on \(x\), a checker control state, an input position \(i \le \left\lvert x\right\rvert \), and a counter \(r\), the length-checker configuration retains the tape bank of \(c\) together with the output of \(c\) captured on the extra tape, has the given control state and its input head past the first \(i\) cells, and places the capture head at displacement \(r - 1\), so that \(r\) countdown cells remain available.
For every input \(x\) and every position \(i \le \left\lvert x\right\rvert \), in a length-checker configuration at input position \(i\) the input head reads the \(i\)-th bit of \(x\) (the blank when \(i = \left\lvert x\right\rvert \)), independently of the saved tape bank.
If the input splits as \(x = p\, b\, d\, s\) with bits \(b, d\), two checker steps from the block-start parse state at position \(\left\lvert p\right\rvert \) continue silently at position \(\left\lvert p\right\rvert +2\) when \(b = d\); halt with verdict \(0\) when \((b,d) = (1,0)\); and enter the suffix length comparison when \((b,d) = (0,1)\). The countdown counter is unchanged in every case.
If the input splits as \(x = p\, b\, s\) with bit \(b\), one checker step from the block-start parse state at position \(\left\lvert p\right\rvert \) records \(b\) in the finite control and advances the input head; the countdown cells and the output are untouched.
For a split \(x = p s\) of the input and a counter \(r\) at most the captured word’s length, from the checker’s block-start parse state at position \(\left\lvert p\right\rvert \) there is a time \(t \le \left\lvert s\right\rvert + 1\) after which the checker has halted with a one-bit output: if \(s\) decodes under the self-delimiting pair code as \((a,b)\), the bit records whether \(\left\lvert b\right\rvert \le r\), and if \(s\) is malformed the bit is \(0\).
Let \(M\) be a machine that, on input \(x\), halts within \(T\) steps with output \(w\). Then within \(T + \left\lvert x\right\rvert + 4\) steps the length checker built over \(M\) reaches, without emitting any physical output, the checker configuration at input position \(0\) in its parse state, retaining a halted bank of \(M\) whose captured output is \(w\) and with all \(\left\lvert w\right\rvert \) countdown cells available.
For a split \(x = p s\) of the input and a counter \(r\) at most the captured word’s length, from the checker’s comparison state at position \(\left\lvert p\right\rvert \) there is a time \(t \le \left\lvert s\right\rvert + 1\) such that after \(t\) steps the checker has halted with output the single bit recording whether \(\left\lvert s\right\rvert \le r\); the empty suffix succeeds even with counter \(0\).
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet, let \(c_0, \dots , c_N\) be configurations of \(M\) on a common input with \(c_N\) halted with empty output, let \(a\) be a Boolean predicate on round indices, and let \(w\) assign a payload word to each index. Assume for each \(j \lt N\) that within \(B\) steps the machine either halts from \(c_j\) with output \(w(j)\) (when \(a(j)\) holds) or reaches \(c_{j+1}\) (when \(a(j)\) fails). Then from \(c_0\) the machine halts within \(N \cdot B\) steps, with output \(w(i)\) for the least accepting index \(i \lt N\) when one exists and empty output otherwise.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet and \(c\) a configuration whose control state is live. If \(M\) has halted \(t\) steps after \(c\), then there is a time \(u\) with \(0 \lt u \le t\) such that \(M\) is not halted at any time strictly before \(u\), is halted at time \(u\), and the configuration at time \(u\) equals the configuration at time \(t\).
Let \(F\) be a bundled finite machine over the Boolean alphabet that computes, in time \(T\), the function sending each input \(x\) to the binary digit word of \(R(|x|)\). Then for every input \(x\), the length of that digit word is at most \(T(|x|)\), since a run can append at most one output symbol per step.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet, let \(c_0, \dots , c_N\) be configurations of \(M\) on a common input with \(c_N\) halted with output the single bit \(\mathrm{false}\), and let \(a\) be a Boolean predicate on round indices. Assume for each \(j \lt N\) that within \(B\) steps the machine either halts from \(c_j\) with output the single bit \(\mathrm{true}\) (when \(a(j)\) holds) or reaches \(c_{j+1}\) (when \(a(j)\) fails). Then from \(c_0\) the machine halts within \(N \cdot B\) steps with output the single bit recording whether \(a(j)\) holds for some \(j \lt N\).
For every clamped input-head position \(p\) among the \(n+2\) positions of an input of length \(n\) and every direction \(m\) (left, stay, or right), the position reached by one native input-head move in direction \(m\) has value at most the value of \(p\) plus one.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet and \(c\) a configuration of \(M\). If after \(t\) steps from \(c\) the machine has not halted, then for every \(u \le t\) the machine has not halted after \(u\) steps either.
Let \(I\) be a binary predicate on Boolean words, let \(F\) be a step function on Boolean words indexed by an input word, and let \(s_0\) assign an initial word to each input. If \(I(x, s_0(x))\) holds for every input \(x\) and \(I(x, s)\) implies \(I(x, F(x, s))\), then \(I(x, F(x, \cdot )^{i}(s_0(x)))\) holds for every input \(x\) and every iteration count \(i\).
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet with two distinguished control states: a start state whose every transition moves the input head left and enters a scan state, and a scan state that keeps moving left while it reads an input symbol and, on reading the left boundary blank, moves right and enters a given destination state. Then from any configuration \(c\) in the start state there is a time \(t\) at most the input-head position of \(c\) plus two after which the machine sits in the destination state with input head at position one and every other component of \(c\) unchanged.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet and \(c\) a configuration of \(M\). If the output after \(t\) steps from \(c\) is empty, then the output after \(u\) steps is empty for every \(u \le t\), since output is append-only.
Let a host machine agree, on a family of embedded control states, with the capture transformation of the stop wrapper of a loop body, and let \(c\) be a live body configuration together with a release bit and a time \(t\) satisfying the liveness and no- early-anchor guards. Then the host run for \(t\) steps from the capture embedding of \(c\) equals the capture embedding of the simulated stop-wrapper endpoint at time \(t\), including any output bit emitted by a halting transition.
Let \(c\) be a live configuration of the loop body, let \(r\) be a release bit, and let \(t\) be a time such that the body stays live before \(t\) and, except possibly for a released step at time zero, does not visit the anchor before \(t\). Then the stop wrapper run for \(t\) steps from the embedding of \(c\) with release bit \(r\) equals the embedding of the body run at time \(t\), with the release bit consumed after the first action and the flag set to \(\mathrm{true}\) exactly when the body has halted at time \(t\).
Let \(c\) be a configuration of the loop body in a live control state \(q\), and suppose it is not the case that \(q\) is the anchor with the release bit clear. Then one step of the stop wrapper from the embedding of \(c\) equals the embedding of the body’s one-step successor of \(c\), with the release bit cleared and the flag set to \(\mathrm{true}\) exactly when that body step is a halting transition.
Let \(c\) be a configuration of the loop body whose control state is the anchor and whose release bit is clear. One step of the stop wrapper from the embedding of \(c\) halts the wrapper, writes \(\mathrm{false}\) on the flag cell, and changes no other component of the embedded configuration.
For a configuration \(c\) of the loop body, a release bit, and an optional flag bit, the configuration of the stop wrapper whose body tapes carry the data of \(c\), whose extra flag tape stores the flag bit in its origin cell with head at the origin, and whose control state is the control state of \(c\) paired with the release bit.
For a body configuration \(c\), a release bit, an optional flag bit, a counter word \(w\), and a fuel configuration, the configuration of the extended stopped-body source whose leading tapes embed \(c\) with its flag data, whose counter tape is the buffer tape of \(w\) with head at the origin, and whose fuel block carries the fuel configuration’s tapes and heads.
Given a loop body with anchor and a fuel machine, the multi-tape machine on the combined tape count whose transitions are those of the body’s stop wrapper acting on the leading tapes, while the counter tape and the fuel residue tapes are inactive.
Let \(c\) be a configuration of the body’s stop wrapper, and let arbitrary counter and fuel tapes and heads be given. Running the extended stopped-body source for \(t\) steps from the left-block embedding of \(c\) over those inactive tapes equals the left-block embedding of the stop wrapper’s run at time \(t\) over the same inactive tapes.
Given a bundled finite machine (the loop body) and a distinguished anchor control state, the wrapper machine with one extra one-cell flag tape that executes the body but halts silently at the next entry into the anchor, recording on the flag tape whether the stop was an anchor return (\(\mathrm{false}\)) or a genuine body halt (\(\mathrm{true}\)). A set release bit forces one body action even at the anchor, and every successor action clears it; the body’s output is retained in full.
Let \(u\) be a Boolean word and \(j\) the length of its maximal \(\mathrm{false}\) prefix. Then \(2j + 2 \le 2|u| + 2\), and running the one-tape counter decrement machine for exactly \(2j + 2\) steps from its borrow state at cell zero over the buffer tape of \(u\) ends in the return state indexed by the decrement’s success flag, at cell zero, holding the decremented word, with nothing emitted.
Let \(u\) be a Boolean word, \(b\) a flag bit, and \(j \le |u|\). Starting the one-tape counter decrement machine in its rewind state with flag \(b\) at cell \(j - 1\) over the buffer tape of \(u\), after exactly \(j + 1\) steps it is in the return state indexed by \(b\), at cell zero, with the word \(u\) unchanged.
Let \(p\) and \(u\) be Boolean words and let \(j\) be the length of the maximal \(\mathrm{false}\) prefix of \(u\). Starting the one-tape counter decrement machine in its borrow state at cell \(|p|\) of the buffer tape holding \(p\) followed by \(u\), after exactly \(j + 1\) steps it is in the rewind state carrying the decrement’s success flag, at cell \(|p| + j - 1\), with tape contents \(p\) followed by the decremented word.
Let \(p\) and \(b\) be Boolean words. One step of the one-tape counter decrement machine from its borrow state, at cell \(|p|\) of the buffer tape holding \(p\) followed by \(b\), behaves as follows: on an empty \(b\) it enters the rewind state with failure flag one cell to the left without writing; on a leading \(\mathrm{true}\) bit it writes \(\mathrm{false}\), moves left, and enters the rewind state with success flag; on a leading \(\mathrm{false}\) bit it writes \(\mathrm{true}\), moves right, and stays in the borrow state.
For all Boolean words \(p\) and \(b\), the buffer tape holding the concatenation of \(p\) and \(b\) (cells \(0, \dots \) hold the word, every other cell is blank) reads, at cell \(|p|\), the first bit of \(b\), or blank when \(b\) is empty.
For all Boolean words \(p\) and \(b\) and bits \(u\) and \(v\): updating cell \(|p|\) of the buffer tape that holds \(p\) followed by \(u\) followed by \(b\), writing the bit \(v\) there, produces exactly the buffer tape that holds \(p\) followed by \(v\) followed by \(b\).
For a startup marker, a body configuration \(c\), a release bit, an optional flag bit, a counter word, and a fuel configuration, the loop-host configuration obtained by capture- embedding the padded stopped-body configuration, with the return phase chosen according to the startup marker.
For every body call of the loop host, built on a body configuration \(c\) with flag bit, counter word \(w\), and fuel residue, the call equals the controller frame over itself whose flag track stores the flag bit at the origin, whose counter track is the buffer tape of \(w\) with head at zero, and whose payload track is the buffer tape of the output of \(c\) with head at that output’s length, with empty physical output.
Let \(c\) be a body configuration with counter word \(w\) and fuel residue, and let new startup, release, flag, counter-word, and control-state parameters be given. The controller frame built over the original call with the new call’s control state, the new flag at the origin, and the buffer tape of the new counter word equals the call built on \(c\) with the replaced control state and the new parameters.
Let a controller frame of the loop host be given, together with a controller action specified by an input direction, an optional flag write, counter and payload write-and- move pairs, an optional emitted bit, and a successor state. Applying the action to the frame yields the frame with the successor state, the moved input head, the three tracks updated by their optional writes at their current heads, the counter and payload heads moved as instructed, and the emitted bit appended to the output.
For every head direction \(d\), optional emitted bit, and successor state, the loop- controller action that only moves the payload head by \(d\) and emits that bit equals the right-block extension (leaving all earlier tapes stationary) of the corresponding one- tape action.
For an input-head direction, an optional flag write, counter and payload write-and-move pairs, an optional emitted bit, and a successor control state, the action of the loop host that performs exactly these operations; it touches only the flag, counter, and capture (payload) tapes, and every body and fuel tape is stationary under it.
For all Boolean words \(p\) and \(r\) and every bit \(b\), blanking cell \(|p|\) of the partially cleared copying tape with processed prefix \(p\) and suffix \(b\) followed by \(r\) yields the partially cleared copying tape with processed prefix \(p\) extended by \(b\) and suffix \(r\).
For a Boolean word read as a little-endian binary counter, the pair of the decremented word and a success flag: the lowest \(\mathrm{true}\) cell is cleared and all lower cells are set to \(\mathrm{true}\), returning success; if no cell is \(\mathrm{true}\) (underflow), every existing cell is set to \(\mathrm{true}\) and the flag is \(\mathrm{false}\). The word length never changes and the empty word reports underflow.
For every Boolean word \(u\) read as a little-endian counter: if the fixed-width decrement of \(u\) reports success then the value of the decremented word plus one equals the value of \(u\), and if it reports underflow then the value of \(u\) is zero.
For an input word \(x\), an input-head position \(p\), a control state \(q\), a work head position \(z\), and a counter word \(u\), the configuration of the one-tape counter decrement machine whose single work tape is the buffer tape of \(u\), with head at \(z\), control \(q\), input head at \(p\), and empty output.
A bundled finite machine over the Boolean alphabet with one work tape that performs a fixed-width little-endian decrement followed by a rewind to cell zero. Its live control states are a borrow state, rewind states remembering the success flag, and silent waiting return states indexed by that flag; no transition emits physical output.
For a base configuration of the loop host, a control state, an input position, three tape functions for the flag, counter, and payload tracks, two head positions, and an output word, the configuration whose three active controller tracks are given by these data and whose remaining tapes and heads are those of the base configuration.
Let \(c\) be a halted configuration of the fuel machine. The captured fuel endpoint of the loop host equals the controller frame over itself in phase zero whose flag and counter tracks are blank, whose payload track is the buffer tape of the output of \(c\) with payload head at its length, with the input head of \(c\) and empty output.
For a configuration \(c\) of the fuel machine, the configuration of the relocated fuel source that carries the data of \(c\) on the fuel block while the body tapes, the flag tape, and the counter tape are blank with heads at the origin.
Given a loop body and a fuel machine \(F\), the multi-tape machine with the combined tape count whose transitions are those of \(F\) acting on the last block of work tapes, while the body tapes, the flag tape, and the counter tape are left untouched.
Given a loop body with anchor, a fuel machine, and an output mode, the bundled finite machine that runs the fuel machine under capture, rewinds and copies the captured fuel word into a fixed-width binary counter, rewinds the native input, and then repeatedly calls the body between anchor visits: a startup call releases the first round free of charge, each rejecting round decrements the counter in place, and the run ends either at an accepting round – emitting the verdict bit \(\mathrm{true}\) in decision mode or replaying the captured payload in find mode – or at counter underflow, emitting \(\mathrm{false}\) in decision mode and nothing in find mode. Its fourteen controller phases cover fuel rewind, counter copy, head rewinds, startup and round returns, borrow, success and underflow rewinds, exhaustion, and payload replay.
Let \(c\) be a halted body configuration, \(w\) a counter word, and \(f\) a fuel configuration. Then there is a time \(t \le 2\, |o| + 3\), where \(o\) is the output of \(c\), such that from the stopped call on \(c\) with stop flag \(\mathrm{true}\) the loop host at time \(t\) has halted with output \(o\) in find mode and output the single bit \(\mathrm{true}\) in decision mode.
Let \(c\) be a body configuration, \(t\) a time at which the body run from \(c\) sits in the anchor state, with no earlier anchor visit except possibly a released step at time zero, and with the release bit clear whenever \(t = 0\). Then the loop host, run for \(t + 1\) steps from the body call built on \(c\), reaches the stopped call whose body data is the body run at time \(t\) with control removed, release clear, and stop flag \(\mathrm{false}\), with the counter word and fuel residue unchanged.
Let \(c\) be a configuration of the extended stopped-body source on the loop host’s tape layout and \(t\) a time before which that source does not halt. Then the loop host run for \(t\) steps from the capture embedding of \(c\) (with the return phase determined by whether the call is a startup call) equals the capture embedding of the source run at time \(t\).
Let \(w\) be a counter word and \(j\) the length of its maximal \(\mathrm{false}\) prefix. Then \(2j + 2 \le 2|w| + 2\), and running the loop host for exactly \(2j + 2\) steps from its borrow phase on a frame whose counter track is the buffer tape of \(w\) with head at zero ends with the counter track holding the decremented word, head at zero, in the released next-anchor call state on success and in the exhaustion phase on underflow.
Let \(w\) be a counter word, \(b\) a success bit, and \(j \le |w|\). Running the loop host for exactly \(j + 1\) steps from the rewind phase selected by \(b\) on a frame whose counter track is the buffer tape of \(w\) with head at cell \(j - 1\) returns the counter head to zero and dispatches: success releases the next anchor call, underflow enters the exhaustion phase without yet emitting. All inactive residue is retained.
Let \(p\) and \(w\) be Boolean words and \(j\) the length of the maximal \(\mathrm{false}\) prefix of \(w\). Running the loop host for exactly \(j + 1\) steps from its borrow phase on a frame whose counter track holds \(p\) followed by \(w\) with head at \(|p|\) reaches the success- or underflow-rewind phase according to the decrement’s flag, with the counter track holding \(p\) followed by the decremented word and head at \(|p| + j - 1\); all non- counter tracks are preserved.
Let \(p\) and \(r\) be Boolean words. One step of the loop host from its borrow phase on a frame whose counter track is the buffer tape of \(p\) followed by \(r\) with counter head at \(|p|\) behaves as the standalone decrement machine: an empty \(r\) enters the underflow- rewind phase moving left; a leading \(\mathrm{true}\) writes \(\mathrm{false}\) and enters the success-rewind phase moving left; a leading \(\mathrm{false}\) writes \(\mathrm{true}\) and continues the borrow moving right. Only the counter track changes.
Let a loop body with anchor, a fuel machine computing the binary digits of \(R(|x|)\) in time \(T\), an invariant preserved from the initial word by the step function, a bounded anchor-free startup, and a per-round contract (positive duration at most \(T(|x|)\), no early anchor revisit, and accept-with-output or advance-to-stepped-seam) be given, together with an output mode. Then there is a constant \(c\) such that for every input \(x\) the loop host has a round-configuration family: a startup segment of at most \(c (T(|x|) + 1)\) steps reaching round zero, empty output at every round configuration up to \(R(|x|)\), a halted terminal at index \(R(|x|) + 1\) with output empty in find mode and the bit \(\mathrm{false}\) in decision mode, and for each round \(i \le R(|x|)\) a segment of at most \(c (T(|x|) + 1)\) steps that halts with the declared output when round \(i\) accepts and reaches round \(i + 1\) otherwise.
Let \(w\) be a captured payload word and consider the loop host in its replay phase on a frame with empty output whose payload track is the buffer tape of \(w\) with head at zero. After exactly \(|w| + 1\) steps the host has halted and its output is exactly \(w\).
Let \(c\) be a configuration of the relocated fuel source on the loop host’s tape layout and \(t\) a time before which that source does not halt. Then the loop host run for \(t\) steps from the capture embedding of \(c\) equals the capture embedding of the fuel source run at time \(t\).
Let \(p\) be the already-copied prefix and \(r\) the remaining fuel suffix. Starting the loop host in its counter-copy phase on a controller frame whose counter track is the buffer tape of \(p\) and whose payload track is the partially cleared copying tape, with both heads at cell \(|p|\), after exactly \(|r| + 1\) steps the host is in the synchronized- rewind phase with the counter track holding the full fuel word, the payload track blank, and both heads one cell left of the word’s right end.
Let \(w\) be the copied counter word and \(j \le |w|\). Starting the loop host in its synchronized-rewind phase on a controller frame whose counter track is the buffer tape of \(w\) and whose payload track is blank, with both heads at cell \(j - 1\), after exactly \(j + 1\) steps the host is in the input-rewind phase with both heads at cell zero and every other frame component unchanged.
Let \(w\) be a stored word and \(j \le |w|\). Starting the loop host in its fuel-rewind phase on a controller frame whose payload track is the buffer tape of \(w\) with payload head at cell \(j - 1\), after exactly \(j + 1\) steps the host is in the counter-copy phase with the payload head at cell zero and every other frame component unchanged.
Let \(w\) be the captured fuel word. Starting the loop host in controller phase zero on a frame whose counter track is blank and whose payload track is the buffer tape of \(w\) with payload head at cell \(|w|\), after exactly \(3|w| + 4\) steps the host is in the input-rewind phase with the counter track holding \(w\), the payload track blank, both heads at cell zero, and all inactive data unchanged. The empty word is included.
Let \(c\) be a body configuration and \(t \gt 0\) a time such that the body stays live before \(t\), does not visit the anchor strictly between times zero and \(t\), and is halted at time \(t\). Then the loop host, run for \(t\) steps from the active released call on \(c\), reaches the stopped call on the body run at time \(t\) with release clear and stop flag \(\mathrm{true}\), the complete body output captured, including any bit emitted by the halting transition.
Let \(c\) be a configuration of the loop host in controller phase four. Then there is a time \(t\) at most the input-head position of \(c\) plus two after which the host sits in the startup body-call state with input head at position one and every other component of \(c\) unchanged.
Let \(w\) be a captured payload word and \(j \le |w|\). Running the loop host for exactly \(j + 1\) steps from its payload-rewind phase on a frame whose payload track is the buffer tape of \(w\) with payload head at cell \(j - 1\) reaches the replay phase with the payload head at cell zero and every other frame component unchanged.
Let the fuel machine compute, in time \(T\), the function sending each input \(x\) to the binary digit word of \(R(|x|)\). Then for every input \(x\) there are a halted fuel configuration \(c\) with output exactly that digit word and a time \(t \le 5\, T(|x|) + 7\) such that the loop host, run for \(t\) steps from its initial configuration on \(x\), reaches the prepared startup configuration built from \(c\).
Let \(c\) be a halted body configuration with empty output, \(w\) a counter word, and \(f\) a fuel configuration. Then there is a time \(t \le 2|w| + 4\) such that, from the stopped call on \(c\) with stop flag \(\mathrm{false}\): if the fixed-width decrement of \(w\) succeeds, the loop host at time \(t\) is at the released next-anchor call with the decremented counter word; otherwise the host at time \(t\) has halted with output empty in find mode and the single bit \(\mathrm{false}\) in decision mode.
Let \(c\) be a body configuration, \(w\) a counter word, and \(f\) a fuel configuration. One step of the loop host from the startup call on \(c\) with control removed, release clear, and flag \(\mathrm{false}\) yields the active (non-startup) call on \(c\) with control the anchor, release set, and no flag, changing no body, counter, or fuel data.
Let \(w\) be a captured word, let an output prefix and arbitrary completed body and fuel tapes with their head positions be given, and consider the loop host in its replay phase with the payload head at cell zero of the buffer tape of \(w\). After running the host for exactly \(|w| + 1\) steps, the host is halted, its output is extended by exactly \(w\), and the stored word, all inactive tapes, and all inactive heads are unchanged.
Let \(s\) be a round word, \(w\) a counter word, \(f\) a fuel configuration, and \(t \gt 0\) a time such that the body, started at the clean anchor seam carrying \(s\), does not revisit the anchor strictly between times zero and \(t\) and at time \(t\) either halts with a declared payload (acceptance) or sits at the clean seam carrying the stepped word. Then there is a time \(v \le 3t + 2|w| + 5\) such that from the released call on the seam of \(s\): acceptance halts the host with the payload (find mode) or the bit \(\mathrm{true}\) (decision mode); a rejection with successful decrement reaches the released call on the next seam with the decremented counter; and a rejection with underflow halts with empty output (find mode) or the bit \(\mathrm{false}\) (decision mode).
Let \(s\) be a word, \(f\) a fuel configuration, and \(t\) a time such that the body, run from its initial configuration, first visits the anchor at time \(t\), arriving at the clean seam that carries \(s\) on work tape zero with all other work tapes blank. Then the loop host, run for \(t + 2\) steps from the prepared startup configuration, reaches the active released call on that seam with counter word the output of \(f\) and fuel residue \(f\).
For a halted fuel configuration \(c\), the loop-host configuration whose control is the startup body-call state, whose input head is at one, whose flag and payload tracks are blank, whose counter track holds the output word of \(c\) with both active heads at zero, whose output is empty, and which retains the completed fuel residue.
For every halted fuel configuration \(f\), the prepared startup configuration of the loop host equals the startup body call built on the body’s initial configuration with release clear, no flag, counter word the output of \(f\), and fuel residue \(f\).
Let \(p\) and \(r\) be Boolean words and let an output prefix be given. After running the one-tape replay machine for exactly \(|r| + 1\) steps from its live state at cell \(|p|\) of the buffer tape holding \(p\) followed by \(r\), the machine is halted with the head at cell \(|p| + |r|\), the stored word is unchanged, and the output is extended by exactly \(r\) in order.
Let \(p\) and \(r\) be Boolean words. One step of the one-tape replay machine from its live state at cell \(|p|\) of the buffer tape holding \(p\) followed by \(r\) either halts without writing when \(r\) is empty, or, when \(r\) starts with a bit \(b\), appends \(b\) to the output and advances the head by one, leaving the stored word unchanged.
For an input word, an input-head position \(p\), an optional control state, a head position \(z\), a stored word \(w\), and an output prefix, the configuration of the one-tape replay machine whose work tape is the buffer tape of \(w\) with head at \(z\), input head at \(p\), and the given output.
For a machine \(M\), an action of the threaded-map controller built over \(M\): it moves the input head and the final (capture) work head by the given directions, optionally emits one output bit, and passes to the given control state, writing no tape cell.
Applying an administrative action of the threaded-map controller moves the input position and the capture-head displacement by the specified directions, installs the designated control state, and appends the optional emission to the output; no work-tape cell changes.
From the controller’s capture-rewind state with the capture head at displacement \(j-1\), for any \(j\) at most the captured word’s length, exactly \(j+1\) silent steps return the capture head to displacement \(0\) and enter the input-rewind phase, with the input position unchanged.
A configuration of the threaded-map controller that preserves the completed tape bank of the simulated machine and its captured word, and records the controller’s state, the input position, the capture-head displacement, and the output written so far.
In a threaded-map configuration whose input head has passed \(i \le \left\lvert x\right\rvert \) cells, the symbol read is the \(i\)-th bit of \(x\) (the blank when \(i = \left\lvert x\right\rvert \)), independently of the preserved bank.
If the input splits as \(x = p\, b\, d\, s\) with bits \(b, d\), two steps of the threaded-map controller from its block-start parse state at position \(\left\lvert p\right\rvert \) continue parsing when \(b = d\), halt when \((b,d) = (1,0)\), and pass to the post-separator state when \((b,d) = (0,1)\); both bits are emitted exactly when the controller is in its emitting pass. The captured word and its head are preserved.
If the input splits as \(x = p\, b\, s\) with bit \(b\), one step of the threaded-map controller from its block-start parse state at position \(\left\lvert p\right\rvert \) records \(b\) in the finite control and advances the input head, emitting \(b\) exactly when the controller is in its emitting (replay) pass.
For every \(j\) at most the captured word’s length and every prior output \(o\), running the controller for exactly \(j\) steps from its payload-replay state with the capture head at displacement \(0\) yields the payload-replay configuration with the capture head at displacement \(j\) and output \(o\) followed by the first \(j\) captured bits; the bank is preserved.
If the input splits as \(x = p\) followed by the pair encoding of \((a,b)\), then from the controller’s emitting parse state at position \(\left\lvert p\right\rvert \) with prior output \(o\), exactly \(2\left\lvert a\right\rvert + 2\) steps emit the doubled word \(a\) and the separator bits \(0,1\), and enter the captured-payload replay state with the capture head still at displacement \(0\).
Let \(M\) be a machine that, on input \(x\), halts within \(T\) steps with output \(w\). Then within \(T + \left\lvert w\right\rvert + \left\lvert x\right\rvert + 5\) steps the threaded-map controller built over \(M\) reaches, without emitting any physical output, its silent validation state at input position \(0\) with the capture head at \(0\), retaining a halted bank of \(M\) whose captured output is \(w\).
For a split \(x = p s\) of the input, there is a time \(t \le \left\lvert s\right\rvert + 1\) from the controller’s silent validation state at position \(\left\lvert p\right\rvert \) such that: if \(s\) is a valid word of the self-delimiting pair code, after \(t\) steps the controller is at its post-validation state with nothing emitted; and if \(s\) is malformed, after \(t\) steps the controller has halted with empty output.
Let \(p \in \{ 0,\dots ,n+1\} \) be an input-head position for an input of length \(n\). The clamped left move sends \(p\) to the position whose value is \(p - 1\) in truncated natural subtraction; in particular the left boundary position \(0\) is fixed.
Let \(\Gamma \) be a finite alphabet with decidable equality, let \(M\) be a multi-tape Turing machine over \(\Gamma \), and suppose \(M\) computes a function \(f \colon \Gamma ^* \to \Gamma ^*\) within a time bound \(T \colon \mathbb {N}\to \mathbb {N}\). Then there exist a finite alphabet \(\Gamma '\) with decidable equality, an injection \(e \colon \Gamma \hookrightarrow \Gamma '\), a machine \(M'\) over \(\Gamma '\), and a constant \(c \in \mathbb {N}\) such that: in every initialized run of \(M'\) no work head ever visits a negative cell, \(M'\) has the same number of work tapes as \(M\), and \(M'\) computes \(f\) through the embedding \(e\) within the time bound \(n \mapsto c \cdot (T(n) + 1)\).
Let \(M\) be a multi-tape Turing machine over a finite alphabet \(\Gamma \). The predicate asserts that \(M\) uses its work tapes unidirectionally: for every input string over \(\Gamma \), every number of steps \(t \in \mathbb {N}\), and every work tape of \(M\), the position of that work tape’s head after running \(M\) for \(t\) steps from its initial configuration on the input is nonnegative.
For every finite Turing machine \(M\) over an alphabet \(\Sigma \) and all words \(x\) and \(w\) over \(\Sigma \), the machine \(M\) does not compute the output \(w\) from the input \(x\) within \(0\) steps: after zero steps the machine is still in its initial, non-halted state.
A multitape Turing machine \(M\) over an alphabet \(\Gamma \) is oblivious if its head positions are a function of the input length and the time only: for all inputs \(x, y \in \Gamma ^*\) with \(\left\lvert x\right\rvert = \left\lvert y\right\rvert \) and every time \(t \in \mathbb {N}\), the configuration reached after \(t\) steps of \(M\) started on \(x\) has the same input-head position (read as a natural number) and the same tuple of work-tape head positions as the configuration reached after \(t\) steps of \(M\) started on \(y\).
In-model analogue of [AB09, Claim 1.6]. Let \(\Gamma \) be a finite alphabet and let \(M\) be a Turing machine over \(\Gamma \) (read-only input tape, write-only output tape, any number of work tapes) that computes a function \(f\) on words over \(\Gamma \) within time \(T : \mathbb {N} \to \mathbb {N}\), i.e. on every input \(x\) the run of \(M\) has halted after \(T(\left\lvert x\right\rvert )\) steps with output \(f(x)\). Then there exist a finite alphabet \(\Gamma '\), an embedding \(e : \Gamma \hookrightarrow \Gamma '\), a machine \(M'\) over \(\Gamma '\) with exactly one work tape, and a constant \(c \in \mathbb {N}\) such that, on every symbol-wise encoded input \(e(x)\), the run of \(M'\) has halted after \(c \cdot (T(\left\lvert x\right\rvert ) + 1)^2\) steps with output the symbol-wise encoding \(e(f(x))\).
Let \(M\) be a Turing machine over the binary alphabet (read-only input tape, write-only output tape, any number of work tapes) that computes a function \(f\) on binary words within time \(T : \mathbb {N} \to \mathbb {N}\), i.e. on every input \(x\) the run of \(M\) has halted after \(T(\left\lvert x\right\rvert )\) steps with output \(f(x)\). Then there exist a machine \(M'\) over the binary alphabet with exactly one work tape and a constant \(c \in \mathbb {N}\) such that \(M'\) computes \(f\) within time \(n \mapsto c \cdot (T(n) + 1)^2\). This combines the one-work-tape simulation ([AB09, Claim 1.6]) with alphabet reduction ([AB09, Claim 1.5]).
Let \(M\) compute the total function \(g\) within time \(T\). Then the length checker built over \(M\) computes, within \(T(n) + 2n + 5\) steps on inputs of length \(n\), the one-bit function that on input \(x\) outputs whether \(\left\lvert b\right\rvert \le \left\lvert g(x)\right\rvert \) when \(x\) decodes under the self-delimiting pair code as \((a,b)\), and outputs \(0\) when \(x\) is malformed.
Given a machine \(M\), the captured-length checker over \(M\) is a machine with one additional work tape that first simulates \(M\) on the input while capturing the output of \(M\) on the extra tape, rewinds the input head to the start, parses the input’s doubled prefix under the self-delimiting pair grammar, and after the separator compares the length of the remaining suffix against the length of the captured word, emitting a single Boolean verdict; only the final comparison or a rejection emits output.
A five-state machine without work tapes that transforms its input \(x\) into the self-delimiting pair code of \((x,x)\): on a first pass it emits every input bit twice, then it emits the separator bit \(0\), rewinds to the start, emits the separator bit \(1\), and copies \(x\) once more.
On every input \(x\) the pair extractor halts within \(5(\left\lvert x\right\rvert +1)\) steps; if \(x\) decodes under the self-delimiting pair code as \((a,b)\), the output is \(a\) (when the first flag is set) followed by \(b\) (when the second flag is set), and if \(x\) is not a valid pair code the output is empty.
A one-work-tape machine, parameterized by two Boolean flags, that parses an input in the self-delimiting pair code of \((a,b)\): it buffers the decoded first component \(a\) on its work tape while validating the doubled prefix, and after the separator it rewinds and replays the buffer, emitting \(a\) when the first flag is set, then scans the suffix, emitting \(b\) when the second flag is set; invalid inputs make it halt silently.
Let \(M\) compute the total function \(f\) within time \(T\). Then the threaded-map controller built over \(M\) computes, within \(4(T(n) + n + 3)\) steps on inputs of length \(n\), the function sending a word \(x\) that decodes under the self-delimiting pair code with first component \(a\) to the pair encoding of \((a, f(x))\), and any malformed word to the empty word.
Given a machine \(M\), the threaded-map controller over \(M\) is a machine with one extra work tape that first simulates \(M\) on the input while capturing the output of \(M\) on the extra tape, rewinds the capture head and the input, silently validates the input’s membership in the self-delimiting pair code, and only on success rewinds the input once more and emits the input’s doubled prefix and separator followed by the captured word; malformed inputs halt without any emission.
If the input splits as \(x = p\, b\, c\, r\) with bits \(b, c\), then two steps of the validity tester from its block state at position \(\left\lvert p\right\rvert \) continue silently to the block state at position \(\left\lvert p\right\rvert +2\) when \(b = c\), and otherwise halt at position \(\left\lvert p\right\rvert +2\) with single-bit output that is \(1\) exactly when \((b,c) = (0,1)\).
For any split \(x = p s\) of the input, there is a time \(t \le \left\lvert s\right\rvert +1\) such that after \(t\) steps from the block state at position \(\left\lvert p\right\rvert \) with empty output, the validity tester has halted and its output is the single bit recording whether \(s\) is a valid self-delimiting pair code.
A machine without work tapes that checks membership in the self-delimiting pair code: it scans its input in aligned two-bit blocks, remembering the first bit of the current block in its finite control, continues silently while the two bits of a block agree, and halts with a one-bit verdict at the first disagreement or boundary blank; the verdict is \(1\) exactly when the disagreeing block is \(01\).
For every \(j \le \left\lvert x\right\rvert \), after \(j\) steps from its initial configuration on input \(x\) the stripping machine holds the first \(j\) bits of \(x\) on its work tape with both heads advanced by \(j\), and has emitted nothing.
From the stripping machine’s rewind state with retained buffer \(a\) and work head at displacement \(\left\lvert a\right\rvert -1\), exactly \(2(\left\lvert a\right\rvert +1)\) steps reach a halted configuration whose output is exactly \(a\).
From the stripping machine’s replay state with buffer \(a\) and work head at displacement \(0\), for any \(j \le \left\lvert a\right\rvert \), exactly \(j\) steps move the work head to displacement \(j\) and emit the first \(j\) bits of \(a\), preserving the buffer.
From the stripping machine’s rewind state with buffer \(a\) and work head at displacement \(j-1\), for any \(j \le \left\lvert a\right\rvert \), exactly \(j+1\) steps return the work head to displacement \(0\) and enter the replay state, emitting nothing.
From the stripping machine’s reverse state with buffer \(w\) and work head on the final cell, there is a time \(t \le 3(\left\lvert w\right\rvert +1)\) after which the machine has halted; its output is the prefix of \(w\) strictly before the last \(1\) bit of \(w\), and is empty when \(w\) contains no \(1\).
A four-state, one-work-tape machine that copies its input to the work tape, then erases the trailing run of zeros and the last \(1\) from the right, rewinds, and replays the retained prefix to the output; an input containing no \(1\) halts silently during the reverse scan.
Let \(c\) be a configuration of a \(k\)-work-tape machine and \(j \in \mathbb {Z}\). Transducing the block encoding coordinate \(j\) (with unrecorded left-neighbour flags) by the forward-sweep visit rule, starting from the control table associated with \(j\), yields exactly the control table associated with \(j+1\) together with the same block with left-neighbour head flags recorded.
Let \(c\) be a configuration of a \(k\)-work-tape machine, \(j \in \mathbb {Z}\), and \(n \in \mathbb {N}\). The forward sweep over the concatenation of the \(n\) unread blocks at coordinates \(j, \dots , j+n-1\), started from the control table at \(j\), ends in the control table at \(j+n\) and produces the corresponding blocks with left-neighbour head flags recorded.
The control table that the forward sweep holds immediately before reading the block at coordinate \(j\) in a configuration \(c\): for each tape \(i\) it contains the symbol currently scanned by head \(i\) if that head lies at a coordinate strictly less than \(j\) (and nothing otherwise), together with the flag indicating whether head \(i\) is at coordinate \(j-1\).
The visit rule of the forward sweep. Its control state is a table assigning to each tape index an optional symbol and a Boolean flag. On visiting a sweep cell for tape \(i\), the table entry for \(i\) becomes the cell’s payload if the cell’s head flag is set (otherwise the recorded symbol is unchanged), paired with that head flag; the cell keeps its payload and head flag and stores the Boolean previously held in the table, which is the head flag of tape \(i\) in the preceding block.
Let \(M\) be a finite-state Turing machine over the Boolean alphabet with \(k\) work tapes, \(b\) a bit, \(c\) a configuration of \(M\) on an input \(x\), \(a\) an action of \(M\), and \(\rho \) the last bit of the output of \(c\) (empty if that output is empty). Applying the redirected action of \(a\) with register \(\rho \) (the moves and writes of \(a\), no emission, register updated to \(a\)’s emitted bit if any, and on a halt of \(a\) either halting when the updated register equals \(b\) or entering the loop state otherwise) to the redirected image of \(c\) (the tapes and heads of \(c\), empty output, and state determined by \(c\)’s state and the last bit of its output in the same way) gives the redirected image of the configuration obtained by applying \(a\) to \(c\).
Let \(M\) be a finite-state Turing machine over the Boolean alphabet and \(b\) a bit, and consider the halt-redirected machine of \(M\) and \(b\), which simulates \(M\) silently while recording the last bit \(M\) has emitted, halts when \(M\) halts with that bit equal to \(b\), and otherwise enters a loop state \(\ell \). For every configuration \(c\) of the halt-redirected machine on an input \(x\) whose control state is \(\ell \), and every \(t \in \mathbb {N}\), running the halt-redirected machine for \(t\) steps from \(c\) returns exactly \(c\).
Let \(M\) be a finite-state Turing machine over the Boolean alphabet and \(b\) a bit, and consider the halt-redirected machine of \(M\) and \(b\), which simulates \(M\) silently while recording the last bit \(M\) has emitted, halts when \(M\) halts with that bit equal to \(b\), and otherwise enters an idle loop state. For every input \(x\) and every \(t \in \mathbb {N}\), the configuration of the halt-redirected machine after \(t\) steps from its initial configuration on \(x\) is the redirected image of the configuration of \(M\) after \(t\) steps from its initial configuration on \(x\): the same tapes and heads, empty output, and state \((q, \rho )\) if \(M\) is in state \(q\) and the last bit it has emitted is \(\rho \) (empty if none), or, if \(M\) has halted, halted when \(\rho = b\) and in the loop state otherwise.
Let \(M\) be a finite-state Turing machine over the Boolean alphabet and \(b\) a bit, and consider the halt-redirected machine of \(M\) and \(b\), which simulates \(M\) silently while recording the last bit \(M\) has emitted, halts when \(M\) halts with that bit equal to \(b\), and otherwise enters an idle loop state. For a configuration \(c\) of \(M\) on an input \(x\), let its redirected image carry the tapes and heads of \(c\), empty output, and the state \((q, \rho )\) when \(c\) is in state \(q\) and \(\rho \) is the last bit of its output (empty if none), or, when \(c\) is halted, the halting state if \(\rho = b\) and the loop state otherwise. Then for every such \(c\), halted or not, one step of the halt-redirected machine from the redirected image of \(c\) gives the redirected image of one step of \(M\) from \(c\).
Let \(S\) be a set of states, \(b\) a bit, \(a\) an action of a Boolean Turing machine with \(k\) work tapes and state set \(S\), and \(\rho \) either a bit or empty. Let \(\rho '\) be the bit \(a\) emits if there is one, and \(\rho \) otherwise. The redirected action is the action on \(k\) work tapes with state set \((S \times \{ \text{empty}, \mathrm{false}, \mathrm{true}\} ) \sqcup \{ \ell \} \) that has the same input move and work-tape writes and moves as \(a\), emits nothing, and has successor \((s', \rho ')\) if \(a\) has successor \(s'\), while if \(a\) halts it halts when \(\rho ' = b\) and enters the loop state \(\ell \) otherwise.
Let \(M\) be a finite-state Turing machine over the Boolean alphabet and \(b\) a bit, and consider the halt-redirected machine of \(M\) and \(b\), which simulates \(M\) silently while recording the last bit \(M\) has emitted, halts when \(M\) halts with that bit equal to \(b\), and otherwise enters an idle loop state \(\ell \). For a configuration \(c\) of \(M\) on an input \(x\), let \(\rho \) be the last bit of the output of \(c\) (empty if that output is empty); the redirected image of \(c\) is the configuration of the halt-redirected machine on \(x\) with the input head, work tapes and work heads of \(c\), empty output, and control state \((q, \rho )\) if \(c\) is in state \(q\), while if \(c\) is halted it is halted when \(\rho = b\) and in \(\ell \) otherwise.
Let \(S\) be a set of states, \(b\) a bit, \(q\) either a state in \(S\) or the halting marker, and \(\rho \) either a bit or empty. The redirected control state, an optional element of \((S \times \{ \text{empty}, \mathrm{false}, \mathrm{true}\} ) \sqcup \{ \ell \} \), is the simulation state \((s, \rho )\) if \(q = s \in S\); if \(q\) is the halting marker, it is the halting marker when \(\rho = b\) (so never when \(\rho \) is empty) and the loop state \(\ell \) otherwise.
Let \(M\) be a Turing machine over the Boolean alphabet with finitely many states, state set \(Q\), initial state \(q_0\) and \(k\) work tapes, and let \(b\) be a bit. The halt-redirected machine of \(M\) and \(b\) has \(k\) work tapes, control states the pairs \((s, \rho )\) with \(s \in Q\) and \(\rho \) either a bit or empty, together with one loop state \(\ell \), and initial state \((q_0, \text{empty})\). In state \((s, \rho )\) it performs \(M\)’s input move and work-tape writes and moves at \(s\) but emits nothing, updates the register to \(\rho '\), the bit \(M\) emits if there is one and \(\rho \) otherwise, and enters \((s', \rho ')\) if \(M\) moves to state \(s'\), while if \(M\) halts it halts when \(\rho ' = b\) and enters \(\ell \) otherwise; in state \(\ell \) it moves no head, writes and emits nothing, and stays in \(\ell \).
Let \(M\) be a finite-state Turing machine over the Boolean alphabet, \(b\) a bit, \(x\) and \(w\) Boolean words, and \(t \in \mathbb {N}\). Consider the halt-redirected machine of \(M\) and \(b\), which simulates \(M\) without emitting anything while recording the last bit \(M\) has emitted, halts when \(M\) halts with that bit equal to \(b\), and otherwise enters an idle loop state. If \(M\) on input \(x\) halts within \(t\) steps with output \(w\), and \(w\) is nonempty with last bit \(b\), then the halt-redirected machine on input \(x\) halts within \(t\) steps with empty output.
Let \(M\) be a finite-state Turing machine over the Boolean alphabet, \(b\) a bit, \(x\) and \(w\) Boolean words, and \(t \in \mathbb {N}\). Consider the halt-redirected machine of \(M\) and \(b\), which simulates \(M\) without emitting anything while recording the last bit \(M\) has emitted, halts when \(M\) halts with that bit equal to \(b\), and otherwise enters an idle loop state. If \(M\) on input \(x\) halts within \(t\) steps with output \(w\), and either \(w\) is empty or its last bit differs from \(b\), then for every \(u \in \mathbb {N}\) the halt-redirected machine, run for \(u\) steps from its initial configuration on \(x\), is not halted.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet on input \(x\), with a start state \(r\), a scanning state \(s\), and an optional dispatch state \(d\), such that regardless of the symbols read: from \(r\) the machine always moves the input head left and enters \(s\); from \(s\) it moves left and stays in \(s\) when an input symbol is scanned, and moves right and passes to \(d\) when a blank is scanned; and none of these transitions touches a work tape or the output. Then from any configuration \(c\) in state \(r\) there exists a number of steps \(t\) such that after \(t\) steps the configuration agrees with \(c\) in every component except that its state is \(d\) and its input head is at position \(1\). This holds even when \(x\) is empty or the input head of \(c\) starts at a boundary blank.
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet on input \(x\), let \(s\) be a scanning state, and let \(d\) be an optional dispatch state, such that regardless of the work symbols read: in state \(s\) scanning an input symbol, \(M\) moves the input head left and stays in \(s\); in state \(s\) scanning a blank, \(M\) moves the input head right and passes to \(d\); and neither transition touches a work tape or the output. Then from any configuration \(c\) in state \(s\) whose input-head position \(p\) satisfies \(p \le \left\lvert x\right\rvert \), running \(M\) for exactly \(p + 1\) steps yields the configuration that agrees with \(c\) in every component except that its state is \(d\) and its input head is at position \(1\), the position scanning the first input symbol.
Given natural numbers \(k\) and \(l\), a state renaming \(f : S \to S'\), and a one-step action \(a\) of a machine with \(l\) work tapes over the Boolean alphabet, the action of a machine with \(k + l\) work tapes that performs \(a\) on the last \(l\) work tapes, writes nothing and moves no head on the first \(k\) tapes, keeps the input-head movement and the emitted output of \(a\), and renames the successor state along \(f\), preserving halting.
Given a state renaming \(f : S \to S'\), a configuration \(c\) of a machine with \(l\) work tapes over the Boolean alphabet on input \(x\), and arbitrary contents and head positions for \(k\) further tapes, the configuration of a machine with \(k + l\) work tapes on the same input whose state is that of \(c\) renamed along \(f\) (halted exactly when \(c\) is halted), whose input head and output agree with those of \(c\), whose last \(l\) work tapes and heads are those of \(c\), and whose first \(k\) work tapes carry the given inactive contents and head positions, for instance data left behind by an earlier phase.
Let \(f : S \to S'\) be a state renaming, \(a\) a one-step action of a machine with \(l\) work tapes over the Boolean alphabet, \(c\) a configuration of such a machine on input \(x\), and fix contents and head positions for \(k\) inactive extra tapes. Applying the right-block extension of \(a\) along \(f\) to the right-block embedding of \(c\) with the given inactive data yields the right-block embedding, with the same inactive data, of the result of applying \(a\) to \(c\).
Let \(M\) be a machine with \(l\) work tapes and \(M'\) a machine with \(k + l\) work tapes, both over the Boolean alphabet, and let \(f\) be a map from the states of \(M\) to the states of \(M'\) such that on every renamed state, for all scanned symbols, the transition of \(M'\) is the right-block extension along \(f\) of the corresponding transition of \(M\) read off the last \(l\) tapes. Then for every configuration \(c\) of \(M\) on input \(x\), all inactive contents and head positions for the first \(k\) tapes, and every number of steps \(t\), running \(M'\) for \(t\) steps from the right-block embedding of \(c\) equals the right-block embedding of the run of \(M\) for \(t\) steps from \(c\), with the inactive left-block data preserved throughout.
Let \(M\) be a machine with \(l\) work tapes and \(M'\) a machine with \(k + l\) work tapes, both over the Boolean alphabet, and let \(f\) be a map from the states of \(M\) to the states of \(M'\) such that on every renamed state \(f(q)\), for every scanned input symbol and every tuple of scanned work symbols, the transition of \(M'\) is the right-block extension along \(f\) of the transition of \(M\) from \(q\) on that input symbol and the symbols scanned on the last \(l\) tapes. Then for every configuration \(c\) of \(M\) on input \(x\) and all inactive contents and head positions for the first \(k\) tapes — which may hold arbitrary data from an earlier phase — one step of \(M'\) from the right-block embedding of \(c\) equals the right-block embedding of one step of \(M\) from \(c\).
The right guard block of \(k\) sweep cells determined by a read table \(s\): the cell for tape \(i\) carries index \(i\), no payload, a cleared head flag, and a left-neighbour flag equal to the Boolean recorded for tape \(i\) in \(s\) — the head flag of the last scanned block.
Suppose a control state \(q\) of a machine without work tapes emits each scanned bit and moves right, and halts silently on the blank. Then \(\left\lvert x\right\rvert +1\) steps from the scan configuration at position \(0\) with prior output \(o\) reach a halted configuration with output \(o\) followed by all of \(x\).
Suppose a control state \(q\) of a machine without work tapes emits each scanned bit and moves right, and halts silently on the blank. Then for every \(j \le \left\lvert x\right\rvert \), \(j\) steps from the scan configuration at position \(0\) with prior output \(o\) lead to the scan configuration at position \(j\) in state \(q\) with output \(o\) followed by the first \(j\) bits of \(x\).
Suppose a control state \(q\) of a machine without work tapes emits each scanned bit and moves right, and halts silently on the blank. If \(x = p s\) is a split of the input into a prefix \(p\) and suffix \(s\), then \(\left\lvert s\right\rvert +1\) steps from the scan configuration at position \(\left\lvert p\right\rvert \) in state \(q\) with prior output \(o\) reach a halted configuration with output \(o\) followed by \(s\).
If at control state \(q\) and input position \(i \lt \left\lvert x\right\rvert \) the transition of a machine without work tapes moves right, optionally emits a bit \(e\), and passes to state \(q'\), then one step from the scan configuration at \(i\) with output \(o\) yields the scan configuration at \(i+1\) in state \(q'\) with output \(o\) extended by \(e\) when present.
Suppose a control state \(q\) of a machine without work tapes moves right on reading a \(1\), emitting a \(0\) when an emission flag is set and nothing otherwise. If the first \(j \le \left\lvert x\right\rvert \) bits of \(x\) are all \(1\), then \(j\) steps from the scan configuration at position \(0\) in state \(q\) with empty output lead to position \(j\) in state \(q\), with output \(j\) zeros when the flag is set and empty output otherwise.
Let \(M\) be a finite Turing machine over an alphabet \(\Gamma \) (a deterministic multi-tape machine with finitely many states and some number of work tapes), let \(x\) be an input word, and let \(t \in \mathbb {N}\). In the configuration reached by running \(M\) for \(t\) steps from its initial configuration on \(x\): every work-tape head position lies in the interval \([-t, t]\), and on every work tape, every cell at a coordinate \(z\) with \(z \lt -t\) or \(z \gt t\) is blank.
A configuration of the source machine over the empty virtual input in which every work tape holds the unary word of length one more than the candidate’s length, all heads are at displacement \(0\), and the control state and output are given parameters.
If the source machine, started on the prepared bank, reaches the halted bank with output \(o\) after \(T\) steps, then there is a time \(t \le T\) at which the round body, started from the counting-host embedding of the start bank, reaches the embedding of the halted bank with output \(o\); the whole trace avoids the anchor.
For any input \(w\) and candidate \(s\), there is a time \(t \le 2(\left\lvert s\right\rvert +1)\) such that \(t+1\) steps of the round body from its preparation seam reach the counting-host embedding of the prepared source bank at the source’s start state, and the whole trace avoids the anchor.
For any input \(w\) and candidate \(s\), there is a time \(t \le 2\left\lvert s\right\rvert + \left\lvert w\right\rvert + 5\) such that the round body, started from the restoration-phase embedding of the cleanup scan, reaches the canonical anchored-word configuration in the restoration phase’s return state carrying one split-search step applied to \(s\); the whole trace avoids the anchor.
For a configuration \(c\) of the rewind machine in its start state, there is a time \(t \le i + 2\), with \(i\) the current input position, such that the round body run from the rewind-phase embedding of \(c\) (carrying the acceptance bit) reaches the embedding of \(c\) with its input head back at the first cell and in the rewind machine’s return state; the whole trace avoids the anchor and preserves the source bank and output.
Suppose the source machine, started on the prepared bank for candidate \(s\), reaches the halted bank with output \(o\) after \(T\) steps. Then there is a time \(t\) with \(0 \lt t \le T + 5\left\lvert s\right\rvert + 3\left\lvert w\right\rvert + 20\) such that the round body, started from the canonical anchored configuration carrying \(s\) on input \(w\), never revisits the anchor strictly between times \(0\) and \(t\), and at time \(t\): if \(\left\lvert s\right\rvert + \left\lvert o\right\rvert = \left\lvert w\right\rvert \) the body has halted with output the pair encoding of the split of \(w\) at position \(\left\lvert s\right\rvert \); otherwise the body is back at the canonical anchored configuration carrying one split-search step applied to \(s\).
The finite control of the combined split-search round body over a source state type: an anchor state, a preparation phase, a counting phase carrying a source state and an overflow flag, a check phase, an input-rewind phase carrying the acceptance bit, a restoration phase, and an emission phase; only restoration can return to the anchor.
Given a source machine \(M\) and a start state, the machine executing one round of the split search: from the anchor it dispatches to preparation, runs the embedded preparation machine, starts \(M\) on the prepared bank with its emissions counted against the native input, checks whether the consumed length matches the input length, rewinds the input carrying the verdict, and then either runs the split emitter (acceptance) or the rejection cleanup followed by an explicit return to the anchor.
In the counting-host embedding of a source configuration, the conjunction of the negated overflow flag with the test that the native input read is blank holds exactly when the candidate’s length plus the source’s output length equals the input length.
Applying the transformed action to the counting-host embedding of a source configuration yields exactly the counting-host embedding of the source configuration after the original action: one native input cell is consumed per source emission, the candidate is preserved, and all source-bank writes and moves are reproduced.
Suppose the host’s transitions on embedded states agree with the emission-counting transform of the source’s transitions, and the source run from \(c\) reaches a halted configuration \(d\) after \(T\) steps. Then there is a time \(t \le T\) at which the host run from the embedding of \(c\) equals the embedding of \(d\).
For a configuration whose input head is at the saturated position for \(j\) on input \(w\), the disjunction of the flag recording \(\left\lvert w\right\rvert \lt j\) with the test that the symbol read is blank equals the flag recording \(\left\lvert w\right\rvert \lt j + 1\).
Suppose the host machine’s transitions on embedded states agree with the emission-counting transform of the source machine’s transitions, and the source, started from a configuration \(c\) over the empty input, does not halt before time \(t\). Then \(t\) steps of the host from the embedding of \(c\) land exactly on the embedding of the source configuration after \(t\) steps; the host’s physical output stays empty throughout.
The transform that converts an action of a source machine into an action of the counting host: each output emission of the source becomes one rightward move of the native input head, the first work tape (holding the candidate) is untouched, the source’s tape writes and moves are reproduced on the successor-indexed tapes, nothing is emitted, and a finite flag records consumption past the native right boundary.
The configuration of the counting host that embeds a source configuration over the empty virtual input: the candidate \(s\) sits on the first work tape, the source’s work tapes occupy the remaining tapes, and the native input head is at the saturated position for \(\left\lvert s\right\rvert \) plus the length of the source’s output so far.
Let a source machine have a class of terminal control states on which it stalls, and let a host machine agree, on the images of non-terminal states under a state embedding, with the mapped source transitions, where no embedded state equals a distinguished host anchor. If the source run from \(c\) reaches a terminal configuration \(d\) within \(T\) steps, then there is a \(t \le T\) such that the host run from the embedded \(c\) reaches the embedded \(d\) at time \(t\) and the entire host trace up to time \(t\) avoids the anchor.
If a host machine’s transitions on all embedded states equal the mapped transitions of a source machine, then for every start configuration \(c\) and time \(t\), the host run from the embedded \(c\) equals the embedding of the source run from \(c\), including through halting.
Let the candidate’s length be at most \(\left\lvert w\right\rvert \). For every \(j\) at most the candidate’s length, \(2j\) steps of the split emitter from its start emit each of the first \(j\) bits of \(w\) twice and advance both the native head and the candidate counter to \(j\); candidate bits are read only for their presence.
If the candidate’s length \(\ell \) is at most \(\left\lvert w\right\rvert \), then exactly \(\ell + \left\lvert w\right\rvert + 3\) steps of the split emitter from its start produce the self-delimiting pair encoding of the split of \(w\) at position \(\ell \) (the first \(\ell \) bits paired with the rest) and halt; the candidate may contain any bit pattern.
If the native input splits as \(w = p r\), then from the split emitter’s copy state at position \(\left\lvert p\right\rvert \) with prior output \(o\), exactly \(\left\lvert r\right\rvert +1\) steps copy \(r\) to the output and halt, preserving all work tapes; this includes the empty suffix with its final blank-reading halt.
Given an input \(w\), a candidate \(s\), an optional control state, a count \(j\), a counter-head position \(h\), and an output word, the emitter configuration is the configuration of the split emitter with that control state, with its input head at the saturated position for \(j\), with \(s\) on its first work tape and the head there at \(h\), with every scratch tape holding the unary word of length \(\left\lvert s\right\rvert +1\) and its head at \(0\), and with the given output.
A four-state machine with \(k+1\) work tapes that emits the self-delimiting pair encoding of a split of its native input: using the first work tape only as a length counter, it emits each native bit twice while the counter has cells left, then emits the separator bits \(0,1\), then copies the remaining native suffix verbatim; no counter bit is ever emitted.
For all \(C\) and \(e\) and every word \(w\), searching the indices \(0, \dots , \left\lvert w\right\rvert \) for the first \(i\) at which the split-acceptance test succeeds on the \(i\)-th iterate of the split step returns exactly the arithmetic split search at parameters \(C\), \(e\), and \(\left\lvert w\right\rvert \), i.e. the least \(i\) with \(i + C\, (i+1)^{e} = \left\lvert w\right\rvert \) when one exists and nothing otherwise.
The arithmetic split search at parameters \(C\), \(e\), and \(\left\lvert w\right\rvert \) fails if and only if the split-acceptance test rejects the \(i\)-th iterate of the split step on the empty candidate for every \(i \le \left\lvert w\right\rvert \).
Splitting \(w\) at the length of the first accepted orbit candidate and pair encoding the two halves gives the same word as splitting \(w\) at the index returned by the arithmetic split search; when both searches fail, both sides are the empty word.
On the empty input, with every unary loop tape of side \(q \gt 0\) already installed and all heads at \(0\), the unary polynomial generator runs for exactly one more step than the exact nest cost at level \(c+1\) and halts, having emitted \(C\, q^{c+1}\) ones, with every loop head back at \(0\).
For every input \(w\) and candidate \(s\), one step of the preparation machine from the end of its scan yields the prepared configuration in its rewind phase, in which every scratch tape holds the unary word of length \(\left\lvert s\right\rvert +1\) and all work heads are at displacement \(\left\lvert s\right\rvert -1\); the uniform extra cell makes the empty candidate need no special case.
For any input \(w\) and candidate \(s\), there is a time \(t \le 2(\left\lvert s\right\rvert +1)\) such that the preparation machine, started from the canonical anchored configuration carrying \(s\), never visits its return state before time \(t\) and at time \(t\) sits in the prepared configuration with all heads at \(0\).
For every \(j\) at most the candidate’s length, exactly \(j+1\) steps of the preparation machine from the prepared configuration in its rewind phase with all work heads at displacement \(j-1\) yield the prepared configuration in its return phase with every head at displacement \(0\); each scratch tape keeps its extra unary cell.
For every input \(w\) and candidate \(s\), exactly \(2(\left\lvert s\right\rvert +1)\) silent steps of the preparation machine from the canonical anchored configuration whose stored words carry \(s\) yield the prepared configuration, in which every scratch tape holds the unary word of length \(\left\lvert s\right\rvert +1\) and every head is at displacement \(0\).
For every \(j\) at most the candidate’s length, \(j\) steps of the preparation machine from the start of its scan yield the preparation scan configuration for \(j\); the candidate is unchanged and nothing is emitted, and this holds also for candidates longer than the native input.
A configuration of the preparation machine in which every scratch tape holds the unary word of length one more than the candidate’s length, all work heads sit at a common displacement, and the finite flag records whether the candidate is longer than the input.
For \(k\) scratch tapes, a machine with \(k+1\) work tapes that prepares the polynomial loop bank of the split search: it scans the candidate on the first work tape while writing its length in unary onto every scratch tape in parallel, adds one extra side-length cell, rewinds all work heads along the untouched candidate, and stalls in an absorbing return state.
One step of the cleanup machine from the end of its scan phase clears the final extra scratch cell and appends a \(1\) to the candidate exactly when the candidate’s length is at most the input length; the result is the cleaned configuration holding one split-search step applied to the candidate.
For any input \(w\) and candidate \(s\), there is a time \(t\) with \(0 \lt t \le 2\left\lvert s\right\rvert + \left\lvert w\right\rvert + 5\) such that the cleanup machine never visits its return state before time \(t\) and at time \(t\) sits in the canonical anchored return configuration carrying one split-search step applied to \(s\).
For every \(j\) at most the candidate’s length, exactly \(j+1\) steps of the cleanup machine from the cleaned configuration with all work heads at displacement \(j-1\) yield the cleaned configuration of the next phase with every work head at displacement \(0\); no candidate bit is altered.
For any input \(w\) and candidate \(s\), there is a time \(t \le 2\left\lvert s\right\rvert + \left\lvert w\right\rvert + 5\) after which the cleanup machine, started at the beginning of its scan phase, reaches the canonical anchored configuration, in its return state, whose stored words carry one split-search step applied to \(s\); every transition on the way is silent.
For every \(j\) at most the candidate’s length, \(j\) steps of the cleanup machine from the start of its scan phase reach the cleanup scan configuration for \(j\), with nothing emitted; the overflow flag records exactly whether more candidate cells than input cells have been consumed.
For \(k\) scratch tapes, a machine with \(k+1\) work tapes implementing the rejection cleanup of the split search: it scans the candidate word on the first work tape while clearing every unary scratch tape, appends one \(1\) bit to the candidate exactly when the candidate’s length does not exceed the input length, rewinds all work heads along the candidate and then rewinds the input head, and finally stalls in an absorbing return state.
For all constants \(C\) and \(e\) there exist a machine \(M\) and a constant \(c\) computing, within \(c\, (n+1)^{e+2}\) steps on inputs of length \(n\), the function sending \(w\) to the pair encoding of the split of \(w\) at the least \(i\) with \(i + C\, (i+1)^{e} = \left\lvert w\right\rvert \), and to the empty word when no such \(i\) exists.
Let a body machine with a distinguished anchor state satisfy a startup contract (within \(A(n+1)^{e+1}\) steps it reaches the anchored empty-candidate configuration without an earlier anchor visit) and a round contract (from the anchored configuration carrying any candidate of length at most \(n+1\), within \(A(n+1)^{e+1}\) steps and with no interior anchor visit, it either halts with the encoded split when the candidate satisfies the length equation, or returns to the anchor with the stepped candidate). Then there exist a machine \(M\) and a constant \(c\) computing, within \(c\, (n+1)^{e+2}\) steps, the function that sends \(w\) to the pair encoding of the split of \(w\) at the least \(i\) with \(i + C\, (i+1)^{e} = \left\lvert w\right\rvert \), and to the empty word when no such \(i\) exists.
Let a source machine, started on the prepared bank for any candidate \(s\), halt within \(B(\left\lvert s\right\rvert )\) steps with captured output a run of \(C\, (\left\lvert s\right\rvert +1)^{e}\) ones, where \(B(l) \le (C+1+5e)(l+1)^{e} + 1\). Then there exist a machine \(N\) and a constant \(c\) computing, within \(c\, (n+1)^{e+2}\) steps, the function sending \(w\) to the pair encoding of the split of \(w\) at the least \(i\) with \(i + C\, (i+1)^{e} = \left\lvert w\right\rvert \), and to the empty word when no such \(i\) exists.
The fixed-prefix emitter for the word \(1^C\), started on the empty virtual input in the bank for candidate \(s\) (it has no scratch tapes), halts after exactly \(C+1\) steps with captured output \(1^C\), the final step being the blank-reading halt.
For all \(c\) and \(C\) and every candidate \(s\), the unary polynomial generator with \(c+1\) loop levels and \(C\) emissions per innermost visit, started in its outermost loop state on the prepared bank for \(s\), halts after exactly one more step than the exact nest cost at side length \(\left\lvert s\right\rvert +1\), and its captured output is a run of \(C\, (\left\lvert s\right\rvert +1)^{c+1}\) ones.
Given an input \(x\), an optional control state, an input position \(i \le \left\lvert x\right\rvert \), a buffer word \(w\), a head displacement \(h\), and an output word, the stripping-machine configuration is the configuration of the marker-stripping machine with that control state, with its input head past the first \(i\) cells, with \(w\) stored contiguously on its work tape and the work head at \(h\), and with the given output.
Let \(M\) be a machine with \(k \ge 1\) work tapes and let \(c\) be a configuration with a state \(q\) such that every work-tape head of \(c\) is at a coordinate strictly less than \(a+n\) and every work tape of \(c\) is blank at \(a+n\). From the configuration reached after the forward sweep — the right-extension phase carrying \(q\) and the read table for coordinate \(a+n\), with the read-marked encoded blocks for coordinates \(a-1, \dots , a+n-1\) on the tape between boundary markers — running exactly \(k + 1 + (n+2)k + 1\) transitions writes the right guard block, performs the transition of \(M\) from \(q\) on the scanned input symbol and the collected work-tape symbols (moving the input head and emitting output accordingly), sweeps backward applying that action to every block, and reaches the canonical between-steps configuration of the resulting source configuration with a zone of \(n+2\) blocks starting at \(a-1\).
Let \(M\) be a one-work-tape Turing machine over alphabet \(A\), let \(w\) be a word over \(A\) to be written, let \(d\) be a fixed integer displacement, and let \(\mathit{cfg}\) be a family of configurations of \(M\) indexed by a step counter \(i \in \{ 0, \dots , |w|\} \), an integer coordinate \(z\), and two lists \(l\), \(r\) of possibly-blank symbols. Assume the local one-step rule: for every \(i \lt |w|\) and all \(z\), \(l\), \(r\), a single step of \(M\) from \(\mathit{cfg}(i, z, l, r)\) yields \(\mathit{cfg}(i+1,\, z + d,\, w_i \mathbin {::} l,\, \mathrm{tail}(r))\), where \(w_i\) is the \(i\)-th symbol of \(w\), pushed as a nonblank cell. Then for every \(n \le |w|\), running \(M\) for exactly \(n\) steps from \(\mathit{cfg}(0, z, l, r)\) yields \(\mathit{cfg}\) at counter \(n\), coordinate \(z + dn\), first list the first \(n\) symbols of \(w\), reversed, prepended to \(l\), and second list \(r\) with its first \(n\) entries removed. The hypothesis is purely local and carries no premise about the global run.
For the single-tape simulator of a machine \(M\) with \(k\) work tapes and any source state \(q\): from any backward-facing zipper configuration whose control is in the left-extension phase carrying \(q\) with counter \(0\), running exactly \(k\) transitions writes the encoded unmarked blank block in reversed cell order while moving left, ending in the left-extension phase with counter \(k\) and the work head moved left by \(k\) cells; the input head and output are unchanged.
For the single-tape simulator of a machine \(M\) with \(k\) work tapes, a source state \(q\), and a read table \(s\): from any forward-facing zipper configuration whose control is in the right-extension phase carrying \(q\) and \(s\) with counter \(0\), running exactly \(k\) transitions writes the guard block determined by \(s\) while moving right, ending in the right-extension phase with counter \(k\), still carrying \(q\) and \(s\), with the work head advanced by \(k\) cells; the input head and output are unchanged.
Let \(M\) be a machine with \(k\) work tapes over a finite alphabet \(\Gamma \) and let \(x\) be an input word. Running the single-tape simulator for exactly \(2k+2\) transitions from its initial configuration on the symbol-wise encoded input reaches the canonical between-steps configuration representing the initial configuration of \(M\) on \(x\), with a zone of one block at coordinate \(0\) and work head at absolute cell \(-1\).
For the single-tape simulator of a machine \(M\) with \(k\) work tapes: from any forward-facing zipper configuration whose control is in the initialization phase with counter \(0\), running exactly \(k\) transitions writes, one cell per step while moving right, the encoded marked blank block over the next \(k\) cells, ending in the initialization phase with counter \(k\) and with the work head advanced by \(k\) cells; the input head, output, and remaining tape contents are unchanged.
Let \(M\) be a machine with \(k \ge 1\) work tapes and let \(c\) be a configuration of \(M\) in a non-halted state \(q\) such that every work-tape head of \(c\) is at a coordinate at least \(a\) and every work tape of \(c\) is blank at \(a-1\). From the canonical between-steps configuration for \(c\) with a zone of \(n\) blocks starting at \(a\), running the simulator for exactly \(k + 1 + (n+1)k + 1\) transitions extends the zone by a blank block at \(a-1\), performs the forward sweep over all \(n+1\) blocks — collecting the scanned symbol of every source tape into the control and recording left-neighbour head flags in the cells — and ends in the right-extension phase carrying \(q\) and the completed read table for coordinate \(a+n\), with the read-marked encoded blocks enclosed by boundary markers on the tape and the input head and encoded output preserved.
Let \(M\) be a one-work-tape Turing machine over alphabet \(A\) with states \(S\), and let controller states \(R\) and controller cells \(C\) be encoded into \(S\) and \(A\) by maps \(\mathit{state}\) and \(\mathit{symbol}\). Suppose a local rule \(\mathit{visit} \colon R \times C \to R \times C\) is implemented by \(M\)’s transition table: whenever \(M\) is in state \(\mathit{state}(s)\) and its work head reads \(\mathit{symbol}(a)\) (for any input symbol), its transition is the local sweep action that writes the encoding of the rewritten cell, moves right, and enters the encoding of the new controller state. Then, starting from the forward-sweep configuration with state \(\mathit{state}(s)\), frontier \(z\), left stack \(l\), and right stack consisting of the encoded word \(\mathit{as}\) followed by \(r\), running \(M\) for exactly \(|\mathit{as}|\) steps reaches the forward-sweep configuration whose state encodes the final controller state of the transduction of \(\mathit{as}\), whose frontier is \(z + |\mathit{as}|\), and whose left stack is the encoded rewritten word, reversed, prepended to \(l\), with right stack \(r\); the input-head position and the output word are unchanged.
Let \(M\) be a one-work-tape Turing machine over alphabet \(A\) with states \(S\), and let controller states \(R\) and controller cells \(C\) be encoded into \(S\) and \(A\) by maps \(\mathit{state}\) and \(\mathit{symbol}\). Suppose a local rule \(\mathit{visit} \colon R \times C \to R \times C\) is implemented by \(M\)’s transition table: whenever \(M\) is in state \(\mathit{state}(s)\) and its work head reads \(\mathit{symbol}(a)\) (for any input symbol), its transition is the local sweep action that writes the encoding of the rewritten cell, moves left, and enters the encoding of the new controller state. Then, starting from the backward-sweep configuration with state \(\mathit{state}(s)\), frontier \(z\), left stack \(l\), and right stack consisting of the encoded word \(\mathit{as}\) followed by \(r\), running \(M\) for exactly \(|\mathit{as}|\) steps reaches the backward-sweep configuration whose state encodes the final controller state of the transduction of \(\mathit{as}\), whose frontier is \(z - |\mathit{as}|\), and whose left stack is the encoded rewritten word, reversed, prepended to \(l\), with right stack \(r\); the input-head position and the output word are unchanged.
Let \(M\) be a machine with \(k \ge 1\) work tapes over a finite alphabet, let \(x\) be an input word, and suppose the run of \(M\) from its initial configuration on \(x\) has not halted at any time \(t \lt \tau \). Then for every \(t \le \tau \), running the single-tape simulator for exactly \(\sigma _k(t)\) transitions (the exact simulation time after \(t\) source steps) from its initial configuration on the symbol-wise encoded input reaches the canonical between-steps configuration representing the configuration of \(M\) after \(t\) steps, with a zone of \(2t+1\) blocks starting at coordinate \(-t\) and work head at absolute cell \(-tk - 1\).
Let \(M\) be a machine with \(k \ge 1\) work tapes and let \(c\) be a configuration in a non-halted state \(q\) whose work-tape heads all lie in the interval \([a, a+n)\) and whose work tapes are blank at \(a-1\) and at \(a+n\). Then, from the canonical between-steps configuration for \(c\) with a zone of \(n\) blocks starting at \(a\), exactly \((2n+5)k + 4\) transitions of the single-tape simulator reach the canonical between-steps configuration for the configuration obtained by one step of \(M\) from \(c\), with a zone of \(n+2\) blocks starting at \(a-1\) and the absolute work-head position lowered by \(k\).
Consider a one-work-tape configuration in forward-facing zipper form: work head at cell \(z\), a list \(r\) occupying cells \(z, z+1, \dots \), a list \(l\) occupying cells \(z-1, z-2, \dots \), an input-head position, and an accumulated output. Applying an action that writes a symbol \(b\) at the current cell, moves the work head left, moves the input head in a given direction, optionally emits one output symbol, and enters a given optional state produces the backward-facing zipper at \(z-1\): cells \(z, z+1, \dots \) hold \(b\) followed by the tail of \(r\), cells \(z-1, z-2, \dots \) hold \(l\), the input head has moved in the given direction, and the output is extended by the emitted symbol.
Consider a one-work-tape configuration in backward-facing zipper form: work head at cell \(z\), a list \(r\) occupying cells \(z, z-1, \dots \), and a list \(l\) occupying cells \(z+1, z+2, \dots \). Applying the sweep action that writes a symbol \(b\) at the current cell, moves the work head right, leaves the input head and output unchanged, and enters a state \(q'\) yields the forward-facing zipper at \(z+1\): cells \(z, z-1, \dots \) hold \(b\) followed by the tail of \(r\), and cells \(z+1, z+2, \dots \) hold \(l\).
For a machine state \(q\), a possibly-blank symbol \(a\), and a direction \(d\), this defines the one-work-tape machine action that overwrites the scanned work-tape cell with \(a\), moves the work head one step in direction \(d\), leaves the input head stationary, appends nothing to the output, and passes to state \(q\).
The enlarged alphabet of the single-tape simulator over a source alphabet \(\Gamma \) with \(k\) tapes: a symbol is either a data symbol from \(\Gamma \) (used unchanged for input and output), or an internal sweep cell, or a distinguished boundary marker. Formally it is the disjoint union of \(\Gamma \) with the set of sweep cells extended by one extra element.
For an alphabet \(\Gamma \) and \(k \in \mathbb {N}\), a sweep cell is a tuple consisting of a tape index \(i \in \{ 0, \dots , k-1\} \), an optional payload symbol from \(\Gamma \) (absence meaning blank), a Boolean head flag, and a second Boolean flag in which the forward sweep records the head flag of the corresponding cell of the preceding block.
Given an optional machine state \(q\), an input-head position \(p\), an integer frontier \(z\), two zipper stacks \(l\) and \(r\) of possibly-blank symbols, and an output word, this defines the configuration of a one-work-tape Turing machine whose work tape is the two-stack zipper with frontier \(z\), left stack \(l\), and right stack \(r\), whose work head sits at \(z\), and whose input-head position and output are the given ones.
Consider the forward-sweep configuration of a one-work-tape machine with (optional) state \(q\), input-head position \(p\), frontier \(z\), left stack \(l\), right stack \(a \mathbin {::} r\), and output word \(\mathit{out}\). Applying to it the local sweep action that writes the possibly-blank symbol \(b\), moves the work head right, and enters state \(q'\) produces exactly the forward-sweep configuration with state \(q'\), the same input-head position \(p\) and output \(\mathit{out}\), frontier \(z + 1\), left stack \(b \mathbin {::} l\), and right stack \(r\).
Consider the forward-sweep configuration of a one-work-tape machine with (optional) state \(q\), input-head position \(p\), frontier \(z\), left stack \(l\), an arbitrary (possibly empty) right stack \(r\), and output word \(\mathit{out}\). Applying the local sweep action that writes the possibly-blank symbol \(b\), moves the work head right, and enters state \(q'\) produces the forward-sweep configuration with state \(q'\), the same input-head position and output, frontier \(z + 1\), left stack \(b \mathbin {::} l\), and right stack the tail of \(r\); when \(r\) is empty, the cell consumed is a blank beyond the stored zone.
Given a local rule \(\mathit{visit} \colon R \times C \to R \times C\) mapping a control state and a cell to a new state and a rewritten cell, and an initial state \(s\), this defines the left-to-right transduction of a word over \(C\): the word is processed cell by cell, each cell being rewritten by \(\mathit{visit}\) while the control state is threaded through, and the result records both the final control state and the fully rewritten word.
Let \(\mathit{visit} \colon R \times C \to R \times C\) be a local rule, \(s\) an initial control state, and \(\mathit{as}\), \(\mathit{bs}\) two words over \(C\). The finite-state transduction of the concatenation \(\mathit{as} \mathbin {+\! \! +} \mathit{bs}\) starting from \(s\) is obtained by transducing \(\mathit{as}\) from \(s\), then transducing \(\mathit{bs}\) from the resulting control state: the final state is that of the second transduction, and the rewritten word is the concatenation of the two rewritten words.
Let \(c\) be a configuration of a \(k\)-work-tape machine on input \(x\), and let \(d\) be a configuration of any machine over the enlarged sweep alphabet on the symbol-wise encoded input whose input head is at the transported position of \(c\)’s input head. Then decoding the input symbol scanned by \(d\) yields exactly the input symbol scanned by \(c\).
For a one-work-tape machine, the action that keeps the input head stationary, writes nothing on the work tape (preserving the scanned cell) while moving its head in a given direction \(d\), emits no output, and makes a given optional state the next state.
For an input word \(x\) over \(\Gamma \), the transport of an input-head position on \(x\) (one of \(\left\lvert x\right\rvert +2\) positions, including the two endmarkers) to the numerically identical position on the symbol-wise encoded input, which has the same length.
Given an optional machine state \(q\), an input-head position \(p\), an integer frontier \(z\), two zipper stacks \(l\) and \(r\) of possibly-blank symbols, and an output word, this defines the configuration of a one-work-tape Turing machine whose work head sits at \(z\) and whose work tape is the coordinate reflection of the two-stack zipper: the cell at coordinate \(w\) is the zipper tape with frontier \(-z\), left stack \(l\), and right stack \(r\), read at \(-w\). Thus \(r\) stores the cells at the frontier and to its left (in scanning order toward decreasing coordinates) and \(l\) stores the cells strictly to the right.
Consider the backward-sweep configuration of a one-work-tape machine with (optional) state \(q\), input-head position \(p\), frontier \(z\), left stack \(l\), right stack \(a \mathbin {::} r\), and output word \(\mathit{out}\). Applying to it the local sweep action that writes the possibly-blank symbol \(b\), moves the work head left, and enters state \(q'\) produces exactly the backward-sweep configuration with state \(q'\), the same input-head position \(p\) and output \(\mathit{out}\), frontier \(z - 1\), left stack \(b \mathbin {::} l\), and right stack \(r\).
Consider the backward-sweep configuration of a one-work-tape machine with (optional) state \(q\), input-head position \(p\), frontier \(z\), left stack \(l\), an arbitrary (possibly empty) right stack \(r\), and output word \(\mathit{out}\). Applying the local sweep action that writes the possibly-blank symbol \(b\), moves the work head left, and enters state \(q'\) produces the backward-sweep configuration with state \(q'\), the same input-head position and output, frontier \(z - 1\), left stack \(b \mathbin {::} l\), and right stack the tail of \(r\); when \(r\) is empty, the cell consumed is a blank beyond the stored zone.
In a backward-facing zipper configuration (work head at cell \(z\), a list \(r\) occupying cells \(z, z-1, \dots \), a list \(l\) occupying cells \(z+1, z+2, \dots \)), applying the sweep action that writes a symbol \(b\) at the current cell without moving, leaves the input head and output unchanged, and enters a state \(q'\) replaces the first entry of \(r\) by \(b\) and the control state by \(q'\), leaving everything else fixed.
The canonical simulator configuration representing a source configuration \(c\) between simulated steps, parametrised by a start coordinate \(j\), a block count \(n\), and an absolute cell \(z\). The control is in the left-extension phase carrying the state of \(c\) (halted when \(c\) is halted); the input head is at the transported position of \(c\)’s input head; the output holds the symbol-wise encoding of \(c\)’s output. The work-tape head scans a boundary marker at cell \(z\); the cells above it hold the encoded blocks for coordinates \(j, j+1, \dots , j+n-1\) in ascending order, closed by a second boundary marker, and all remaining cells are blank.
The finite state space of the single-tape simulator of a \(k\)-work-tape machine with state set \(S\) over \(\Gamma \). Its phases are: initialization, carrying a counter in \(\{ 0, \dots , k\} \); a return-to-origin walk; left extension of the zone, carrying a source state and a counter; the forward reading sweep, carrying a source state and a table assigning to each tape an optional symbol and a Boolean; right extension, additionally carrying a counter; and the backward writing sweep, carrying a source state, the scanned input symbol, the collected reads, and a table of Booleans. No unbounded tape coordinate enters the state, so the state space is finite.
Given an integer frontier coordinate \(z\) and two finite lists \(l\) and \(r\) of possibly-blank symbols over an alphabet \(A\), this defines the tape content \(\mathbb {Z} \to A \cup \{ \text{blank}\} \) that stores the cells strictly to the left of \(z\) in \(l\) (nearest cell first) and the cells at \(z\) and to its right in \(r\) (in order); every position beyond the extent of either list reads as blank.
Let \(z\) be an integer and \(l\) a finite list of possibly-blank symbols over an alphabet \(A\). The two-stack zipper tape with frontier \(z\), left stack \(l\), and empty right stack equals the zipper tape with the same frontier and left stack whose right stack is the single blank cell.
Let \(z\) be an integer and let \(l\) and \(r\) be finite lists of possibly-blank symbols over an alphabet \(A\). The two-stack zipper tape with frontier \(z\), left stack \(l\), and right stack \(r\) reads, at position \(z\), exactly the first entry of \(r\) — or blank when \(r\) is empty.
Let \(z\) be an integer, let \(l\) and \(r\) be finite lists of possibly-blank symbols over an alphabet \(A\), and let \(a\), \(b\) be possibly-blank symbols. Overwriting position \(z\) with \(b\) in the two-stack zipper tape with frontier \(z\), left stack \(l\), and right stack \(a \mathbin {::} r\) yields exactly the zipper tape with frontier \(z + 1\), left stack \(b \mathbin {::} l\), and right stack \(r\): the written cell transfers from the right stack to the left stack.
Let \(z\) be an integer and let \(l\) and \(r\) be finite lists of possibly-blank symbols over an alphabet \(A\). The two-stack zipper tape with frontier \(z\), left stack \(l\), and right stack \(r\) coincides, as a function on \(\mathbb {Z}\), with the coordinate reflection \(w \mapsto \) (zipper tape with frontier \(-(z - 1)\), left stack \(r\), and right stack \(l\), read at \(-w\)): reversing the scanning direction exchanges the two stacks.
Consider the two-way infinite partial tape assembled from a position \(z \in \mathbb {Z}\), a list \(r\) of optional symbols occupying cells \(z, z+1, \dots \), and a list \(l\) occupying cells \(z-1, z-2, \dots \) (all unlisted cells blank). Overwriting the cell at \(z\) with a value \(b\) yields the tape assembled from \(z\), \(l\), and the list whose first entry is \(b\) followed by the tail of \(r\).
The single-work-tape simulator of a machine \(M\) with \(k\) work tapes over a finite alphabet \(\Gamma \): a machine over the enlarged sweep alphabet, with one work tape, whose finite control realises the phases of the simulation. Initialization writes the marked origin block and its two boundary markers. Each simulated step extends the stored zone by a blank block on the left, sweeps forward reading the marked payloads while recording left-neighbour head flags in the cells, extends the zone on the right, performs the source transition on the input head and output, and sweeps backward rewriting the cells, carrying right-neighbour head flags in its control.
Given a family of values indexed by \(\{ 0,\dots ,k-1\} \) (the left block), a single value (the buffer), and a family of values indexed by \(\{ 0,\dots ,l-1\} \) (the right block), the family indexed by the \(k + (1 + l)\) tape indices that returns the left values on the first \(k\) indices, the buffer value on the next index, and the right values on the last \(l\) indices.
The block of \(k\) sweep cells encoding coordinate \(j \in \mathbb {Z}\) of a configuration \(c\) of a \(k\)-work-tape machine: the cell for tape \(i\) carries the symbol of tape \(i\) at coordinate \(j\) and the flag indicating whether head \(i\) is at \(j\). When a Boolean parameter is set, the cell additionally records whether head \(i\) is at \(j-1\); otherwise that field is false.
Let \(c\) be a configuration of a \(k\)-work-tape machine and \(j \in \mathbb {Z}\) a coordinate such that no work-tape head of \(c\) is at \(j\) and every work tape of \(c\) is blank at \(j\). Then the unread block encoding coordinate \(j\) equals the blank unmarked block.
For any family of blocks indexed by \(\mathbb {Z}\) and any \(j \in \mathbb {Z}\) and \(n, m \in \mathbb {N}\), the zone of \(n+m\) blocks starting at \(j\) is the concatenation of the zone of \(n\) blocks starting at \(j\) with the zone of \(m\) blocks starting at \(j+n\).
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet with \(k_D\), \(k_1\), \(k_2\) work tapes, consider the timed conditional controller for them, and for a bit \(\beta \) consider the two-branch union machine that runs \(M_1\) or \(M_2\) on disjoint blocks of \(k_1\) and \(k_2\) work tapes and starts in the initial state of \(M_1\) if \(\beta \) is \(\mathrm{true}\) and of \(M_2\) otherwise. Let \(c\) be a configuration of the union machine on an input \(x\), \(\tau \) and \(h\) contents and head positions for \(k_D\) tapes, and \(t \in \mathbb {N}\). Then running the controller for \(t\) steps from the branch-phase configuration of \((c, \tau , h, \beta )\) (the state, input head, output and work tapes of \(c\) on the branch block, \(\tau \) and \(h\) on the decider block, and the bit \(\beta \) in cell \(0\) of the capture tape with that head at cell \(0\)) yields the branch-phase configuration of \((c', \tau , h, \beta )\), where \(c'\) is the configuration the union machine reaches after \(t\) steps from \(c\).
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet, and consider the timed conditional controller for them, which first runs \(D\) silently on a decider block of tapes while copying \(D\)’s output onto a capture tape. Let \(c\) be a configuration of \(D\) on an input \(x\) and \(t \in \mathbb {N}\) such that \(D\) run from \(c\) is not halted at any time \(s \lt t\). Then running the controller for \(t\) steps from the decider-phase configuration of \(c\) (state of \(c\), or the back state if \(c\) is halted; \(c\)’s input head, work tapes and heads on the decider block; a blank branch block; \(c\)’s output stored from cell \(0\) of the capture tape with that head just past it; empty output) yields the decider-phase configuration of the configuration \(D\) reaches after \(t\) steps from \(c\).
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet, and consider the timed conditional controller for them, with a decider block, a branch block and a capture tape as work tapes. For every input \(x\), the initial configuration of the controller on \(x\) equals the decider-phase configuration of \(D\)’s initial configuration on \(x\): \(D\)’s initial state, the initial input-head position, every work tape of all three blocks blank with its head at cell \(0\), and empty output.
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet, and consider the timed conditional controller for them, with a decider block, a branch block and a capture tape as work tapes. Let \(c\) be a halted configuration of \(D\) on an input \(x\) whose output is the one-bit word \(\beta \). Then two steps of the controller from the decider-phase configuration of \(c\) (in the back state, with \(c\)’s input head and work tapes on the decider block, a blank branch block, \(\beta \) in cell \(0\) of the capture tape with that head at cell \(1\), and empty output) reach the ready-to-rewind configuration of \(c\) and \(\beta \), which differs only in being in the rewind-start state for \(\beta \) with the capture head at cell \(0\).
Let \(M\) be a multi-tape Turing machine with \(k\) work tapes over the Boolean alphabet on input \(x\), with a start state \(r\), a scanning state \(s\), and an optional dispatch state \(d\), such that regardless of the symbols read: from \(r\) the machine always moves the input head left (clamped at the left boundary blank) and enters \(s\); from \(s\) it moves left and stays in \(s\) when an input symbol is scanned, and moves right and passes to \(d\) when a blank is scanned; and none of these transitions touches a work tape, a work head, or the output. Let \(c\) be a configuration of \(M\) in state \(r\) whose input head is at position \(p \in \{ 0,\dots ,\left\lvert x\right\rvert +1\} \), where position \(0\) is the left boundary blank. Then there exists \(t \le p + 2\) such that running \(M\) for \(t\) steps from \(c\) yields the configuration that agrees with \(c\) in every component except that its state is \(d\) and its input head is at position \(1\).
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet with \(k_D\), \(k_1\), \(k_2\) work tapes, consider the timed conditional controller for them, let \(x\) be a Boolean word, \(\beta \) a bit and \(T \in \mathbb {N}\), and suppose \(D\) on input \(x\) halts within \(T\) steps with output the one-bit word \(\beta \). Then there exist \(a \le 2T + 5\) and contents \(\tau \) and head positions \(h\) for \(k_D\) tapes such that after \(a\) steps from its initial configuration on \(x\) the controller is in the branch-phase configuration built from the initial configuration on \(x\) of the two-branch union machine selecting \(\beta \): in the copy of the initial state of \(M_1\) if \(\beta \) is \(\mathrm{true}\) and of \(M_2\) otherwise, with the initial input-head position, a blank branch block with every head at cell \(0\), the decider block holding \(\tau \) with heads \(h\), the bit \(\beta \) in cell \(0\) of the capture tape with that head at cell \(0\), and empty output.
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet with \(k_D\), \(k_1\), \(k_2\) work tapes, and consider the timed conditional controller for them, whose work tapes form a decider block of \(k_D\) tapes, a branch block of \(k_1 + k_2\) tapes and one capture tape, and whose states include a copy of the states of \(M_1\) and \(M_2\). Let \(c\) be a configuration on an input \(x\) of the two-branch union machine (which runs \(M_1\) on the first \(k_1\) and \(M_2\) on the last \(k_2\) of its \(k_1 + k_2\) work tapes, according to whether its state belongs to \(M_1\) or \(M_2\)), let \(\tau \) and \(h\) be contents and head positions for \(k_D\) tapes, and let \(\beta \) be a bit. The branch-phase configuration is the controller configuration on \(x\) whose state is the copy of \(c\)’s state (halted if \(c\) is halted), whose input head and output are those of \(c\), whose decider block holds \(\tau \) with heads \(h\), whose branch block holds \(c\)’s work tapes and heads, and whose capture tape holds the bit \(\beta \) in cell \(0\), blank elsewhere, with its head at cell \(0\).
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet with \(k_D\), \(k_1\), \(k_2\) work tapes and state sets \(Q_D\), \(Q_1\), \(Q_2\). The timed conditional controller has \(k_D + k_1 + k_2 + 1\) work tapes, split into a decider block of \(k_D\) tapes, a branch block of \(k_1 + k_2\) tapes and one capture tape; its states are those of \(Q_D\), two administrative states (back and read), a rewind-start and a rewind-scan state for each bit \(\beta \), and those of \(Q_1 \sqcup Q_2\); and its initial state is \(D\)’s. In a state of \(D\) it performs \(D\)’s transition on the input and the decider block, writes \(D\)’s emitted bit (if any) on the capture tape and moves that head right, emits nothing, leaves the branch block idle, and enters the back state where \(D\) would halt; the back state moves the capture head one cell left and enters the read state; the read state reads the bit \(\beta \) under the capture head (\(\mathrm{false}\) if blank) and enters rewind-start for \(\beta \); rewind-start moves the input head left and enters rewind-scan; rewind-scan keeps moving the input head left while it reads an input symbol and, on reading the blank boundary, moves right and enters the initial state of \(M_1\) if \(\beta \) is \(\mathrm{true}\) and of \(M_2\) otherwise (all four administrative states writing and emitting nothing and keeping the other heads still); and in a state of \(M_i\) it performs \(M_i\)’s transition, including its emission and halting, on the input and on \(M_i\)’s own \(k_i\) tapes within the branch block, leaving all other work tapes idle.
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet with \(k_D\), \(k_1\), \(k_2\) work tapes, and consider the timed conditional controller for them, whose work tapes form a decider block of \(k_D\) tapes, a branch block of \(k_1 + k_2\) tapes and one capture tape, and whose states include a copy of the states of \(D\) and an administrative back state. For a configuration \(c\) of \(D\) on an input \(x\), with output \(u\), the decider-phase configuration is the controller configuration on \(x\) whose state is the copy of \(c\)’s state, or the back state if \(c\) is halted; whose input head is that of \(c\); whose decider block holds \(c\)’s work tapes and heads; whose branch block is blank with every head at cell \(0\); whose capture tape holds \(u\) in cells \(0, \dots , |u|-1\), blank elsewhere, with its head at cell \(|u|\); and whose output is empty.
Let \(D\) be a finite-state Turing machine over the Boolean alphabet with \(k\) work tapes, state set \(Q\) and initial state \(q_0\), and let \(m \in \mathbb {N}\). The padded machine has \(k + m\) work tapes, state set \(Q\) and initial state \(q_0\); in state \(q\), reading input symbol \(\sigma \) and work-tape symbols \(w_0, \dots , w_{k+m-1}\), it performs \(D\)’s action at \((q, \sigma , w_0, \dots , w_{k-1})\) on the input head and the first \(k\) work tapes, with the same emission and successor state, while writing nothing on the last \(m\) work tapes and keeping their heads still.
Let \(D\), \(M_1\), \(M_2\) be finite-state Turing machines over the Boolean alphabet, and consider the timed conditional controller for them, whose work tapes form a decider block, a branch block and a capture tape, and whose states include a rewind-start state for each bit. For a configuration \(c\) of \(D\) on an input \(x\) and a bit \(\beta \), the ready-to-rewind configuration is the controller configuration on \(x\) in the rewind-start state for \(\beta \), with the input head of \(c\), the decider block holding \(c\)’s work tapes and heads, the branch block blank with every head at cell \(0\), the capture tape holding the bit \(\beta \) in cell \(0\), blank elsewhere, with its head at cell \(0\), and empty output.
Let \(\Sigma \) be an alphabet and \(M\) a finite deterministic Turing machine over \(\Sigma \) with \(k\) work tapes, a finite state set, and transition function \(\delta \). The associated finite nondeterministic Turing machine has the same number \(k\) of work tapes and the same state set, and as underlying machine the nondeterministic machine with the same initial state whose two transition functions both equal \(\delta \).
Let \(M\) be a Turing machine over an alphabet \(\Gamma \) with no work tapes that computes a function \(f\) on words over \(\Gamma \) within time \(T\), i.e. on every input \(x\) the run of \(M\) has halted after \(T(\left\lvert x\right\rvert )\) steps with output \(f(x)\). Then the machine obtained from \(M\) by adding one unused work tape also computes \(f\) within time \(T\).
Let \(M\) be a Turing machine over an alphabet \(\Gamma \) with no work tapes, and let \(M'\) be the machine obtained from \(M\) by adding one unused work tape. For every configuration \(c\) of \(M\) on an input \(x\), one transition of \(M'\) applied to the embedding of \(c\) (same state, input position, and output; blank added tape with head at the origin) equals the embedding of the configuration obtained by one transition of \(M\) from \(c\); this also holds when \(c\) is halted.
For a Turing machine \(M\) over an alphabet \(\Gamma \) and a configuration \(c\) of \(M\) on an input \(x\), this produces the configuration of a one-work-tape machine with the same state, input-head position, and output as \(c\), whose single work tape is entirely blank with its head at the origin.
Let \(M\) be a Turing machine over an alphabet \(\Gamma \), with a read-only input tape, a write-only output tape, and no work tapes. This constructs a machine with the same states and exactly one work tape: each transition performs the input-head movement, output emission, and state change of the corresponding transition of \(M\), while writing nothing on the added work tape and keeping its head stationary.
Given a Boolean boundary tag \(b\), an optional scanned symbol, and a requested movement \(m \in \{ -1, 0, +1\} \), the movement actually performed by a buffer head simulating an input head: when the scanned symbol is blank and the requested movement points outward — left with tag false, marking the left boundary, or right with tag true, marking the right boundary — the head stays put; in every other case it moves by \(m\).
Let \(c\) be a configuration of a machine over the Boolean alphabet on input \(w\), let \(p \in \{ 0,\dots ,\left\lvert w\right\rvert +1\} \) be its input-head position, let \(b\) be a Boolean tag valid for \(p\) (forced false at the left boundary \(p = 0\) and true at the right boundary \(p = \left\lvert w\right\rvert + 1\)), and let \(m \in \{ -1, 0, +1\} \) be a requested movement. Write \(p'\) for the result of the native clamped move of \(p\) by \(m\). Then, as integers, \(p - 1\) plus the virtual movement determined by \(b\), the scanned input symbol of \(c\), and \(m\) equals \(p' - 1\); and the updated tag — false after a left move, true after a right move, unchanged when stationary — is again valid for \(p'\). This holds for every word \(w\), including the empty word.
Given a Boolean tag \(b\) and a performed movement \(m \in \{ -1, 0, +1\} \), the tag after the move: false after a left move, true after a right move, and \(b\) itself after a stationary move, so that a suppressed outward move retains its boundary tag and repeated outward attempts remain clamped.
For a position \(p \in \{ 0,\dots ,n+1\} \) of a virtual input head over an input of length \(n\) and a Boolean tag \(b\), the predicate asserting that \(b\) correctly marks the boundaries: if \(p = 0\) then \(b\) is false, and if \(p = n + 1\) then \(b\) is true. At interior positions both tag values are admissible.
Let \(c\) be a configuration of a \(k\)-work-tape machine, \(\alpha \) an action of that machine, and \(j \in \mathbb {Z}\). Folding the backward-sweep visit rule for \(\alpha \) over the reversed block encoding coordinate \(j\) (with recorded left-neighbour flags), starting from the table of head flags at coordinate \(j+1\), yields the table of head flags at coordinate \(j\) together with the reversed block encoding coordinate \(j\) of the configuration obtained by applying \(\alpha \) to \(c\).
Let \(c\) be a configuration of a \(k\)-work-tape machine, \(\alpha \) an action of that machine, \(j \in \mathbb {Z}\), and \(n \in \mathbb {N}\). The backward sweep for \(\alpha \) over the reversed concatenation of the \(n\) read-marked blocks at coordinates \(j, \dots , j+n-1\), started from the table of head flags at coordinate \(j+n\), ends in the table of head flags at \(j\) and produces the reversed unread blocks of the configuration obtained by applying \(\alpha \) to \(c\).
The visit rule of the backward sweep, parametrised by an action \(\alpha \) of the simulated \(k\)-work-tape machine. Its control state is a table of Booleans, one per tape (the head flags seen in the block to the right). On visiting a sweep cell for tape \(i\): the new table records the cell’s head flag; the payload is overwritten by the symbol \(\alpha \) writes on tape \(i\) when this cell carries the head (and kept otherwise); the new head flag is taken from the right-neighbour table, from the current cell, or from the stored left-neighbour flag according as \(\alpha \) moves head \(i\) left, keeps it in place, or moves it right; and the stored left-neighbour flag is cleared.
Let \(w = w_0 w_1 \cdots w_{n-1}\) be a bit string, let \(c \lt n\), and let \(p \in \{ 0, 1\} \). Then the doubling \(\mathrm{dbl}(w)\) (each bit of \(w\) written twice in succession, positions indexed from \(0\)) has a symbol at position \(2c + p\), and that symbol is \(w_c\).
Fixed-width little-endian binary increment with overflow, defined recursively on Boolean words (least significant bit first): the empty word overflows (result none); a word \(\mathrm{false} \cdot r\) is sent to \(\mathrm{true} \cdot r\); and a word \(\mathrm{true} \cdot r\) is sent to \(\mathrm{false} \cdot r'\) where \(r'\) is the increment of \(r\), overflowing whenever the increment of \(r\) overflows. Thus a word of length \(\ell \) is mapped to the binary successor of the same length \(\ell \), and the result is none exactly when the word consists entirely of \(\mathrm{true}\).
Let \(M\) be a deterministic Turing machine with \(k\) work tapes over the Boolean alphabet, with start state \(q_0\), and let \(x\) be a Boolean input word. Then the initial configuration of \(M\) on \(x\) equals the seam configuration \(\mathrm{seam}(q_0, w)\) with \(w_i = \varepsilon \) the empty word for every work tape \(i\); that is, the configuration in state \(q_0\) with the input head on the first input cell, every work tape entirely blank with its head at cell \(0\), and empty output.
Let a configuration of a \(k\)-tape machine on an input string be given, and suppose its input head position equals \(1 + p\) for some natural number \(p\) strictly less than the length of the input. Then the symbol under the input head is precisely the \(p\)-th symbol (zero-indexed) of the input string.
For all bit strings \(x\) and \(y\), the pair encoding \(\langle x, y\rangle \) (each bit of \(x\) written twice, then the separator bits \(0\) then \(1\), then \(y\) verbatim) has length \(\left\lvert \langle x, y\rangle \right\rvert = 2\left\lvert x\right\rvert + 2 + \left\lvert y\right\rvert \).
Let \(M\) be a multi-tape Turing machine over the Boolean alphabet, let \(c_0, \dots , c_N\) be configurations of \(M\) on a common input, all with empty output, and let \(a\) be a Boolean predicate on round indices. Assume that from \(c_N\) the machine halts within \(B\) steps with output the single bit \(\mathrm{false}\), and that for each \(j \lt N\), within \(B\) steps the machine either halts from \(c_j\) with output the single bit \(\mathrm{true}\) (when \(a(j)\) holds) or reaches \(c_{j+1}\) (when \(a(j)\) fails). Then from \(c_0\) the machine halts within \((N+1) \cdot B\) steps, and its output is the single bit recording whether \(a(j)\) holds for some \(j \lt N\).
The algebraic laws of a representation scheme for machines in code normal form: an encoding map from coded machines to bit strings, a total decoding map from bit strings to coded machines (so every string represents some machine), and the law that the code of a machine followed by any number of padding bits \(1\) decodes back to that machine (so every machine has infinitely many representations).
Given a position among the \(n + 2\) input-head positions (the input of length \(n\) flanked by one blank cell on each side) and a sign \(m \in \{ -1, 0, 1\} \), this operation attempts to move the head by \(m\): the sum is computed in \(\mathbb {Z}\) and clamped to the valid range, so an attempted move past either boundary cell results in no movement.
Let \(p\) be an input-head position on an input of length \(n\) with \(p \ne 0\). Then the clamped move of \(p\) by the sign \(-1\) is exactly the position \(p - 1\): a left move that does not start at the left boundary decrements the position.
Let \(p\) be an input-head position on an input of length \(n\) with \(p \ne n + 1\). Then the clamped move of \(p\) by the sign \(1\) is exactly the position \(p + 1\): a right move that does not start at the right boundary increments the position.
A deterministic Turing machine with a read-only input tape, \(k\) work tapes, and a write-only output tape, over an alphabet of symbols extended by a blank. It consists of an initial state \(q_0\) and a transition function that, given the current state, the (possibly blank) symbol under the input head, and the \(k\) symbols under the work-tape heads, returns a one-step action: a movement of the input head, a write-and-move instruction for each work tape, an optional symbol to append to the output, and a successor state or the decision to halt. Neither the alphabet nor the state set is required to be finite.
A function \(f\) is computable within input-indexed time bound \(t\) and space bound \(s\), with respect to injective encodings of its domain and codomain into binary words, if there exist a number \(k\) of work tapes, a finite state set, and a multi-tape Turing machine over the binary alphabet with those tapes and states that computes \(f\) between the given encodings within the bounds \(t\) and \(s\).
Let \(f\) be a function computable within input-indexed time bound \(t\) and space bound \(s\) with respect to given binary encodings, and let \(t'\) and \(s'\) be bounds with \(t(a) \le t'(a)\) and \(s(a) \le s'(a)\) for every input \(a\). Then \(f\) is computable within the bounds \(t'\) and \(s'\) with respect to the same encodings.
A function \(f\) is computable within time bound \(t\) and space bound \(s\), where \(t, s : \mathbb {N} \to \mathbb {N}\) are functions of the encoded input length, if there is a binary Turing machine with finitely many states that, for every input \(a\), computes the encoding of \(f(a)\) from the encoding of \(a\) in at most \(t(n)\) steps and touching at most \(s(n)\) work-tape cells, where \(n\) is the length of the encoding of \(a\).
Given injective encodings of the domain and codomain into words over the machine’s alphabet, a multi-tape Turing machine \(M\) computes a function \(f\) within time bound \(t\) and space bound \(s\) (both functions of the input) if for every input \(a\) there are \(t' \le t(a)\) and \(s' \le s(a)\) such that \(M\), on the encoding of \(a\), halts within \(t'\) steps with output the encoding of \(f(a)\), touching exactly \(s'\) work-tape cells. The alphabet and state set need not be finite.
Let \(M\) be a multi-tape Turing machine that computes a function \(f\) between given injective encodings within input-indexed time bound \(t\) and space bound \(s\), and let \(t'\) and \(s'\) be bounds with \(t(a) \le t'(a)\) and \(s(a) \le s'(a)\) for every input \(a\). Then \(M\) also computes \(f\) between the same encodings within the bounds \(t'\) and \(s'\).
The proposition that a multi-tape Turing machine \(M\), started in its initial configuration on an input word, is halted after \(t\) steps (hence halts within \(t\) steps, since halting is absorbing), that its output tape then holds exactly the given output word, and that the computation up to step \(t\) touches exactly \(s\) work-tape cells. Neither the alphabet nor the state set is required to be finite.
A set \(L \subseteq \alpha \) is decidable within input-indexed time bound \(t\) and space bound \(s\), with respect to an injective binary encoding of \(\alpha \), if the Boolean indicator function of \(L\) is computable within these bounds, the one-bit answer being encoded as the single-symbol output word.
The Boolean predicate stating that a multi-tape Turing machine \(M\) on a given input halts after exactly \(t\) steps: started from its initial configuration, \(M\) is halted after \(t\) steps but not after \(t - 1\) steps (truncated subtraction). In particular the predicate is never true at \(t = 0\): a machine must perform at least one step to halt.
Let \(M\) be a multi-tape Turing machine and let \(c\) be a configuration that is not halted. If running \(M\) for \(t + 1\) steps from \(c\) returns to \(c\) itself, then the computation of \(M\) from \(c\) never halts: for every number of steps \(t'\), the configuration reached after \(t'\) steps from \(c\) is not halted.
Let \(M\) be a deterministic multi-tape Turing machine, let \(x\) be an input word, and let \(t \in \mathbb {N}\). Then the word written on the output tape after \(M\) runs for \(t\) steps from its initial configuration on \(x\) has length at most \(t\).
Let \(M\) be a deterministic multi-tape Turing machine, let \(c\) be any configuration of \(M\), and let \(t \le t'\) be natural numbers. Then the word on the output tape after running \(M\) for \(t\) steps from \(c\) is a prefix of the word on the output tape after running \(M\) for \(t'\) steps from \(c\).
The symbol, if any, that a multi-tape Turing machine \(M\) appends to its output tape when executing one step from a configuration \(c\): the optional output component of the action prescribed by the transition function for the current state and scanned symbols, and no symbol at all when \(c\) is halted.
Given a multi-tape Turing machine \(M\) with \(k\) work tapes and state set \(S\), and a bijection \(e\) from \(S\) to another state set \(S'\), the machine over \(S'\) obtained by relabelling the states of \(M\) along \(e\): its initial state is \(e(q_0)\), and its transition on a state \(q'\), the scanned input symbol, and the \(k\) scanned work-tape symbols is the action that \(M\) prescribes in the state \(e^{-1}(q')\) for the same scanned symbols, with its successor state renamed along \(e\). A bijection, rather than an arbitrary function, is required here: a map identifying two states with different transitions would leave no well-defined transition table.
Let \(M\) be a multi-tape Turing machine with \(k\) work tapes and state set \(S\), let \(e\) be a bijection from \(S\) to another state set, let \(x\) be an input word, and let \(t \in \mathbb {N}\). Then the configuration reached by the machine obtained by relabelling the states of \(M\) along \(e\) after \(t\) steps from its initial configuration on \(x\) equals the configuration that \(M\) reaches after \(t\) steps from its own initial configuration on \(x\), with its state renamed along \(e\).
Let \(M\) be a multi-tape Turing machine with \(k\) work tapes and state set \(S\), let \(e\) be a bijection from \(S\) to another state set, and let \(c\) be a configuration of \(M\) on a fixed input word. Then one step of the machine obtained by relabelling the states of \(M\) along \(e\), taken from the configuration \(c\) with state renamed along \(e\), equals the configuration obtained by one step of \(M\) from \(c\) with state renamed along \(e\). The equality includes the case of a halted \(c\), where both sides leave the configuration unchanged.
The configuration reached by iterating the one-step function of a multi-tape Turing machine \(t\) times starting from a configuration \(c\). Since halted configurations are fixed points of the step function, a machine that halts stays in its halting configuration.
Let \(M\) and \(M'\) be multi-tape Turing machines over the same alphabet, with \(k\) and \(k'\) work tapes and running on inputs \(x\) and \(x'\) respectively, and let \(f\) map configurations of \(M\) to configurations of \(M'\) so that one step of \(M'\) after \(f\) agrees with \(f\) after one step of \(M\) on every configuration. Then for every configuration \(c\) of \(M\) and every \(n \in \mathbb {N}\), running \(M'\) for \(n\) steps from \(f(c)\) yields the image under \(f\) of the configuration \(M\) reaches in \(n\) steps from \(c\).
Let \(M\) be a multi-tape Turing machine, \(c\) a configuration, and \(\tau \le t\) natural numbers such that the configuration reached by \(M\) after \(\tau \) steps from \(c\) is halted. Then the output after \(t\) steps from \(c\) equals the output after \(\tau \) steps from \(c\).
For every multi-tape Turing machine \(M\), configuration \(c\), and \(t \in \mathbb {N}\), running \(M\) for \(t + 1\) steps from \(c\) equals applying one step of \(M\) to the configuration reached after running \(M\) for \(t\) steps from \(c\).
The total number of work-tape cells touched by the computation of a multi-tape Turing machine with \(k\) work tapes from a configuration \(c\) up to step \(t\): the sum over the \(k\) work tapes of the number of distinct cells visited by each head. This is the space measure used throughout; the read-only input tape and the write-only output tape do not count.
Let \(M\) be a multi-tape Turing machine, \(c\) a configuration, and \(t \in \mathbb {N}\). The total space used by \(M\) from \(c\) up to step \(t\) equals the space touched by the list of configurations that \(M\) traverses in steps \(0, 1, \dots , t\) from \(c\), where the space touched by a list of configurations is the sum over the work tapes of the number of distinct head positions occurring in the list.
For every multi-tape Turing machine \(M\), configuration \(c\), step count \(t\), and work-tape index \(i\), the number of cells touched by the head of tape \(i\) up to step \(t\) is at most the total space used by \(M\) from \(c\) up to step \(t\).
The successor of a configuration \(c\) of a multi-tape Turing machine \(M\): if \(c\) is halted (its state component is empty), the step returns \(c\) unchanged; otherwise it applies to \(c\) the action that the transition function of \(M\) prescribes for the current state, the symbol under the input head, and the symbols under the \(k\) work-tape heads.
For every multi-tape Turing machine \(M\) and configuration \(c\), the output component after one step of \(M\) from \(c\) equals the output of \(c\) followed by the symbol emitted by that step, if any (and equals the output of \(c\) unchanged when no symbol is emitted).
Let \(M\) be a multi-tape Turing machine on an input word \(x\), whose input-head position ranges over \(\{ 0, 1, \dots , |x| + 1\} \), and let \(c\) be a configuration of \(M\) with input-head position \(p\). Then for every \(t \in \mathbb {N}\), the input-head position in the configuration reached by running \(M\) for \(t\) steps from \(c\) is at most \(p + t\).
Let \(M\) be a deterministic multi-tape Turing machine with \(k\) work tapes, initial state \(q_0\), and transition function \(\delta \). The associated nondeterministic machine has the same tapes, states, and initial state \(q_0\), and both of its transition functions \(\delta _0\) and \(\delta _1\) are equal to \(\delta \), so that its behaviour does not depend on the choice bit.
Let \(M\) be a deterministic multi-tape Turing machine with transition function \(\delta \), let \(N\) be the nondeterministic machine with the same initial state whose two transition functions both equal \(\delta \), and let \(x\) be an input word. Then the initial configuration of \(N\) on \(x\) coincides with the initial configuration of \(M\) on \(x\).
Let \(M\) be a deterministic multi-tape Turing machine with transition function \(\delta \), and let \(N\) be the nondeterministic machine with the same initial state whose two transition functions both equal \(\delta \). For every configuration \(c\) on a fixed input and every choice word \(w \in \{ 0,1\} ^*\), the run of \(N\) from \(c\) under \(w\) equals the configuration reached by running \(M\) from \(c\) for \(|w|\) steps.
Let \(M\) be a multi-tape Turing machine with \(k\) work tapes, \(c\) a configuration, and \(i\) one of the \(k\) work-tape indices. Then the position of the \(i\)-th work-tape head after one step of \(M\) from \(c\) differs from its position in \(c\) by at most \(1\) in absolute value.
A nondeterministic Turing machine with a read-only input tape, \(k\) work tapes, and a write-only output tape, over an alphabet of symbols extended by a blank, consists of an initial state \(q_0\) and two transition functions \(\delta _0\) and \(\delta _1\), one for each value of a choice bit. Each \(\delta _b\) maps the current state, the (possibly blank) symbol under the input head, and the \(k\) symbols under the work-tape heads to a one-step action: a movement of the input head, a write-and-move instruction for each work tape, an optional symbol to append to the output, and a successor state or the decision to halt. Neither the alphabet nor the state set is required to be finite.
Let \(M\) be a nondeterministic multi-tape Turing machine, \(x\) an input word, and \(t \in \mathbb {N}\). The machine \(M\) halts on \(x\) within \(t\) steps along every branch if, for every choice word \(w \in \{ 0,1\} ^t\) of length exactly \(t\), the configuration reached by running \(M\) under \(w\) from its initial configuration on \(x\) is halted, that is, has an empty state component.
Let \(M\) be a nondeterministic multi-tape Turing machine, \(x\) an input word, and \(t \le t'\) natural numbers. If for every choice word of length \(t\) the run of \(M\) under it from the initial configuration on \(x\) ends in a halted configuration, then the same holds for every choice word of length \(t'\).
The starting configuration of a nondeterministic multi-tape Turing machine \(M\) on an input word \(x\): the machine is in its initial state \(q_0\), the input head scans the first symbol of \(x\), all work tapes are blank with their heads at cell \(0\), and the output tape is empty.
Let \(M\) be a nondeterministic multi-tape Turing machine and \(c\) a configuration of \(M\) on a fixed input. For a choice word \(w = b_1 b_2 \cdots b_m \in \{ 0,1\} ^*\), the run of \(M\) from \(c\) under \(w\) is the configuration obtained from \(c\) by performing \(m\) steps of \(M\), the \(i\)-th of them under the choice bit \(b_i\). Recursively, the run under the empty word is \(c\), and the run under \(b w\) is the run under \(w\) from the configuration reached by one step of \(M\) from \(c\) under \(b\).
Let \(M\) be a nondeterministic multi-tape Turing machine, \(c\) a configuration of \(M\) on some input, and \(w, w'\) choice words. Running \(M\) from \(c\) under the concatenation \(w w'\) gives the same configuration as first running \(M\) from \(c\) under \(w\) and then running \(M\) under \(w'\) from the configuration so reached.
Let \(M\) be a nondeterministic multi-tape Turing machine, \(c\) a configuration of \(M\) on some input, \(b \in \{ 0,1\} \) a choice bit, and \(w\) a choice word. Running \(M\) from \(c\) under the word \(b w\) gives the same configuration as running \(M\) under \(w\) from the configuration reached by one step of \(M\) from \(c\) under the choice bit \(b\).
Let \(M\) be a nondeterministic multi-tape Turing machine and \(c\) a configuration of \(M\) on some input whose state component is empty (the machine has halted). Then for every choice word \(w\), running \(M\) from \(c\) under \(w\) yields \(c\) itself.
Let \(M\) be a nondeterministic multi-tape Turing machine with transition functions \(\delta _0, \delta _1\), let \(b \in \{ 0,1\} \) be a choice bit, and let \(c\) be a configuration of \(M\) on a fixed input. The step of \(M\) from \(c\) under \(b\) returns \(c\) unchanged if \(c\) is halted (its state component is empty), whatever the value of \(b\); otherwise it applies to \(c\) the action that \(\delta _b\) prescribes for the current state, the symbol under the input head, and the symbols under the \(k\) work-tape heads.
Let \(M\) be a nondeterministic multi-tape Turing machine, \(b \in \{ 0,1\} \) a choice bit, and \(c\) a configuration of \(M\) on some input whose state component is empty (the machine has halted). Then one step of \(M\) from \(c\) under \(b\) leaves \(c\) unchanged.
The fixed two-bit serialization of an optional write action on a binary tape, where a present write may itself write the blank: no write is \(00\), writing the blank is \(01\), writing the bit \(0\) is \(10\), and writing the bit \(1\) is \(11\).
An oracle Turing machine with \(k\) ordinary work tapes over a symbol alphabet \(\Sigma \) and state set \(Q\) consists of an initial state \(q_0\), a designated query state \(q_{\mathrm{query}}\), a designated positive-answer state \(q_{\mathrm{yes}}\), a designated negative-answer state \(q_{\mathrm{no}}\), and a transition function that assigns to every state, scanned input symbol (possibly blank), and family of scanned symbols on the \(k+1\) work tapes (the last of which serves as the query tape) a one-step action on the \(k+1\) work tapes. The transition function governs every state except the query state, whose behavior is instead determined by the oracle. The structure itself does not require the three special states to be distinct; that is imposed separately by well-formedness.
Let \(a\) be a one-step action of a machine with \(k'\) work tapes, let \(c\) be a configuration of such a machine, let \(i\) be a work-tape index, and let \(z\) be an integer cell position different from the position of the \(i\)-th work head in \(c\). Then after applying \(a\) to \(c\), the content of cell \(z\) of the \(i\)-th work tape is unchanged.
An oracle machine \(M\) with oracle \(O\) computes the word \(y\) on input \(x\) within \(t\) steps when, after running \(M\) relative to \(O\) for \(t\) steps from its initial configuration on \(x\), the machine has halted and its output tape holds exactly \(y\). This is a time-only notion; space is not constrained.
Let \(M_0\) be a plain multi-tape machine with \(k\) work tapes, let \(O\) be any oracle language, let \(x\) and \(y\) be words, and let \(t \in \mathbb {N}\). The oracle machine obtained by embedding \(M_0\) computes \(y\) on input \(x\) within \(t\) steps relative to \(O\) if and only if, after running \(M_0\) for \(t\) steps from its initial configuration on \(x\), the machine has halted and its output tape holds exactly \(y\). In particular, the input/output behavior and time bounds of the embedded machine do not depend on the oracle.
The initial configuration of an oracle machine with \(k\) ordinary work tapes on a given input word: the machine is in its initial state, the input head is at the origin, all \(k+1\) work tapes (including the query tape) are blank with their heads at the origin, and the output is empty.
The embedding of a plain multi-tape machine with \(k\) work tapes and state set \(Q\) as an oracle machine over the extended state set \(Q \sqcup \{ 0,1,2\} \): the three fresh states serve as the query, positive-answer, and negative-answer states, the initial state is the image of the original initial state, and on original states the transition function is the original transition (renamed along the left inclusion and extended by an inert query tape), so that the machine never moves into a fresh state and ignores the query tape. From each of the three fresh states the transition table halts immediately.
Given an oracle machine \(M\) with \(k\) ordinary work tapes and state set \(Q\), the plain multi-tape machine with \(k+1\) work tapes over the same state set \(Q\) that simulates \(M\) run with the empty oracle: it has the same initial state, and its transition function agrees with that of \(M\) in every state except the query state, where it instead performs a stationary move (no writes, no head movement, no output) directly into the negative-answer state, since the empty oracle always answers no.
For a configuration of a machine with \(k+1\) work tapes over an alphabet \(\Sigma \), the query string is the word read off the query tape (the last work tape) from cell \(0\) rightward, up to but excluding the first blank cell. If no cell at a nonnegative position of the query tape is blank — possible for an arbitrary configuration, though not for one reachable from an initial configuration — the query string is defined to be the empty word.
Let \(M\) be an oracle machine with \(k\) ordinary work tapes, \(O\) an oracle language, \(x\) an input word, and \(t \in \mathbb {N}\). In the configuration reached by running \(M\) with oracle \(O\) for \(t\) steps from its initial configuration on \(x\), the query string — the contents of the query tape from cell \(0\) up to the first blank cell — has length at most \(t\).
Let \(M_0\) be a plain multi-tape machine with \(k\) work tapes, let \(O\) be any oracle language, let \(c\) be a configuration of \(M_0\), and let \(t \in \mathbb {N}\). Running the oracle machine obtained by embedding \(M_0\) for \(t\) steps relative to \(O\) from the embedding of \(c\) yields the embedding of the configuration obtained by running \(M_0\) for \(t\) steps from \(c\).
Let \(M\) be an oracle machine with \(k\) ordinary work tapes, let \(c\) be a configuration with \(k+1\) work tapes, and let \(t \in \mathbb {N}\). Running the plain machine obtained from \(M\) by the empty-oracle elimination for \(t\) steps from \(c\) yields exactly the configuration obtained by running \(M\) for \(t\) steps from \(c\) relative to the empty oracle.
Let \(M\) be an oracle machine with \(k\) ordinary work tapes, \(O\) an oracle language, \(x\) an input word, and \(t \in \mathbb {N}\). Then in the configuration reached by running \(M\) with oracle \(O\) for \(t\) steps from its initial configuration on \(x\), every work-tape cell at an integer position \(z\) with \(t \le \left\lvert z\right\rvert \) is blank. In particular, the query tape always contains a blank cell at a nonnegative position, so the empty-word fallback in the definition of the query string is never invoked in an initialized run.
Let \(M\) be an oracle machine with \(k\) ordinary work tapes, \(O\) an oracle language, and \(x\) an input word. Then for every \(t \in \mathbb {N}\), in the configuration reached by running \(M\) with oracle \(O\) for \(t\) steps from its initial configuration on \(x\): every one of the \(k+1\) work heads is at a position of absolute value at most \(t\), and every work-tape cell at an integer position \(z\) with \(t \le \left\lvert z\right\rvert \) is blank.
One step of an oracle machine \(M\) with \(k\) ordinary work tapes relative to an oracle language \(O\) over the alphabet \(\Sigma \), acting on a configuration \(c\) with \(k+1\) work tapes. A halted configuration is left fixed. If the current state is the query state, the machine moves to the positive-answer state when the query string of \(c\) belongs to \(O\) and to the negative-answer state otherwise, with all tapes, head positions, and the output left unchanged. In every other state, the machine applies the action prescribed by its transition function for the current state and the scanned input and work-tape symbols, exactly like a plain multi-tape machine.
Let \(M\) be an oracle machine with \(k\) ordinary work tapes and let \(c\) be a configuration whose current state is not the query state of \(M\) (in particular, \(c\) may be halted). Then for any two oracle languages \(O_1\) and \(O_2\), one step of \(M\) from \(c\) relative to \(O_1\) equals one step of \(M\) from \(c\) relative to \(O_2\).
Let \(M_0\) be a plain multi-tape machine with \(k\) work tapes, let \(O\) be any oracle language, and let \(c\) be a configuration of \(M_0\). One step, relative to \(O\), of the oracle machine obtained by embedding \(M_0\), taken from the embedding of \(c\) into the \((k+1)\)-tape oracle model, equals the embedding of the configuration obtained by one step of \(M_0\) from \(c\).
Let \(M\) be an oracle machine with \(k\) ordinary work tapes and let \(c\) be a configuration with \(k+1\) work tapes. One step of the plain machine obtained from \(M\) by the empty-oracle elimination, taken from \(c\), equals one step of \(M\) relative to the empty oracle, taken from \(c\).
An oracle Turing machine with \(k\) ordinary work tapes is well formed when its three special states are pairwise distinct: the query state differs from the positive-answer state, the query state differs from the negative-answer state, and the two answer states differ from each other. The initial state is allowed to coincide with the query state (such a machine submits the empty query on its first step).
Let \(M\) be an oracle machine with \(k\) ordinary work tapes, \(O\) an oracle language, \(c\) a configuration, and \(i\) one of the \(k+1\) work-tape indices. Then the position of the \(i\)-th work head after one step of \(M\) relative to \(O\) differs from its position in \(c\) by at most \(1\) in absolute value.
Let \(M\) be an oracle machine with \(k\) ordinary work tapes, \(O\) an oracle language, \(c\) a configuration, \(i\) one of the \(k+1\) work-tape indices, and \(z\) an integer cell position different from the position of the \(i\)-th work head in \(c\). Then after one step of \(M\) relative to \(O\), the content of cell \(z\) of the \(i\)-th work tape is unchanged.
A partial parser on bit strings that reads aligned two-bit blocks from the left: a block \(00\) contributes the bit \(0\) and a block \(11\) contributes the bit \(1\) to the first component, and the first aligned block \(01\) is the separator, at which point the parser returns the accumulated first component together with the remaining suffix, untouched, as the second component. On strings admitting no such parse the result is undefined.
Consider the aligned pair parser on bit strings, which reads aligned two-bit blocks from the left, treats the blocks \(00\) and \(11\) as data bits \(0\) and \(1\), and stops successfully at the first aligned block \(01\). If the parser fails on a bit string \(z\), then \(z = \mathrm{dbl}(w)\, \tau \) for some bit string \(w\) (with \(\mathrm{dbl}(w)\) denoting \(w\) with every bit written twice) and some tail \(\tau \) that is either empty, a single bit, or begins with the bits \(1\) then \(0\).
For all bit strings \(x\) and \(\alpha \), the aligned two-bit parser applied to the pair encoding \(\langle x, \alpha \rangle \) (each bit of \(x\) doubled, then the separator \(01\), then \(\alpha \) verbatim) succeeds and returns exactly the pair \((x, \alpha )\).
Let \(x\) be a bit string, \(w\) an output string, and \(t \le |x|\). Running the six-state diagonal pairing machine for exactly \(t\) steps from the configuration in its copy state (state \(5\)) with input head at position \(1\) and output \(w\) yields the configuration still in the copy state with input head at position \(t + 1\) and output \(w\) followed by the first \(t\) bits of \(x\).
Let \(x\) be a bit string and \(t \le |x|\). Running the six-state diagonal pairing machine from its initial configuration on input \(x\) for exactly \(2t\) steps yields the configuration in the initial control state with input head at position \(t + 1\) whose output is the first \(t\) bits of \(x\) with each bit doubled.
For the six-state diagonal pairing machine on an input string \(x\), in any configuration whose input head sits at position \(j + 1\) with \(j \lt |x|\), the scanned input symbol is the \(j\)-th bit of \(x\), for every control state and every output string.
Let \(x\) be a bit string, \(w\) an output string, and \(j \le |x|\). Running the six-state diagonal pairing machine for exactly \(j + 1\) steps from the configuration in its rewind state (state \(4\)) with input head at position \(j\) and output \(w\) yields the configuration in its copy state (state \(5\)) with input head at position \(1\) and output still \(w\).
For the six-state diagonal pairing machine on an input string \(x\), in any configuration whose input head sits at the right boundary position \(|x| + 1\), the scanned input symbol is the blank, for every control state and every output string — including on the empty input.
For every bit string \(x\) of length \(n\), running the six-state diagonal pairing machine from its initial configuration on input \(x\) for exactly \(4n + 5\) steps yields a halted configuration with input head at position \(n + 1\) whose output is the diagonal pair encoding \(\langle x, x\rangle \) (each bit of \(x\) doubled, then the separator \(01\), then \(x\) verbatim).
Let \(x\) be a bit string and \(w\) an output string. Running the six-state diagonal pairing machine for exactly \(3\) steps from the configuration in its initial control state with input head at the right boundary position \(|x| + 1\) and output \(w\) yields the configuration in its rewind state (state \(4\)) with input head at position \(|x|\) and output \(w\) followed by the separator pair \(01\).
Consider the six-state diagonal pairing machine on an input string \(x\), in a live configuration with control state \(q\), input-head position \(p\), and output string \(w\), and suppose the symbol scanned by the input head is \(b\) (a bit or the blank). Then one step of the machine yields exactly the configuration prescribed by the transition table entry at \((q, b)\): its successor state (or halting), the head position \(p\) moved by the prescribed direction, and the output \(w\) extended by the emitted bit, if any.
A configuration of the six-state diagonal pairing machine on an input string \(x\), specified by an optional control state among the six (absent when the machine has halted), an input-head position in \(\{ 0, \dots , |x| + 1\} \), and the output string emitted so far. The machine has no work tapes, so the work-tape components of the configuration are vacuous.
A fixed six-state Turing machine over the binary alphabet with no work tapes that transforms an input \(x\) into the diagonal pair encoding \(\langle x, x\rangle \). Its six control states implement, in order: emitting the scanned bit while staying put, emitting it again while moving right, emitting the separator bit \(1\) while staying put, taking one unconditional left move, rewinding the input head to the left end, and copying the input verbatim; the first state’s blank branch emits the separator bit \(0\).
The self-delimiting pairing \(\langle x, \alpha \rangle \) of two bit strings: the first string \(x\) with every bit doubled, then the separator pair \(01\), then the second string \(\alpha \) verbatim. Reading aligned two-bit blocks from the left, \(00\) and \(11\) are data and the first aligned \(01\) is the separator (an aligned \(01\) can never occur inside the doubled region), so the encoding determines both components.
For all bit strings \(x\) and \(y\), the self-delimiting pair encoding \(\langle x, y\rangle \) equals the concatenation
where \(\mathrm{dbl}(x)\) is \(x\) with every bit written twice in succession, followed by the separator bits \(0\) then \(1\), followed by \(y\) verbatim.
Let \(n, n' \in \mathbb {N}\) and let \(u, u'\) be bit strings. Write \(\langle x, y\rangle \) for the pair encoding of bit strings \(x\) and \(y\) (each bit of \(x\) written twice, then the separator bits \(0\) then \(1\), then \(y\) verbatim). If \(\langle 1^n, u\rangle = \langle 1^{n'}, u'\rangle \), where \(1^n\) denotes the string of \(n\) ones, then \(n = n'\) and \(u = u'\).
For a Boolean word \(v\), \(\mathrm{cut}(v)\) is the prefix of \(v\) strictly before its last occurrence of \(\mathrm{true}\): if \(v = u \, \mathrm{true} \, \mathrm{false}^m\) for some word \(u\) and \(m \ge 0\), the result is \(u\), and if \(v\) contains no \(\mathrm{true}\) (in particular if \(v\) is empty), the result is undefined (none).
For a tape count \(k\) and a Boolean word \(s\), the assignment of one word per work tape that places \(s\) on tape \(0\) and the empty word on every other tape. It is the standard tape content of a loop body’s round-boundary (seam) configurations: round state on tape zero, all scratch blank.
Let \(M\) be a coded machine, let \(\mathrm{src}\) be a configuration of \(M\), and let \(a\) be a transition action of \(M\). One step of the (unstopped) four-tape table interpreter from the pending-action evaluation configuration — built over the interpreter checkpoint of \(\mathrm{src}\), carrying the eight decoded bits of \(a\) and the halting flag of its successor, with the serialized table, any table cursor \(p\), and the unary tape of the successor index installed — equals exactly the interpreter checkpoint of the configuration obtained by applying \(a\) to \(\mathrm{src}\), with table cursor \(p\); this includes the correct virtual-input head movement, optional write, emission, and the halting case.
From the borrow phase with a given carry flag and the clock word laid out on the clock lane’s zipper, the timed interpreter processes the word in exactly its length: the final control carries the resulting underflow flag, the processed stack holds the borrowed word, and neither the four simulation lanes nor the buffered output is touched.
For every \(j\) at most the clock word’s length, from the clock-rewind phase with the clock head at \(j-1\), exactly \(j+1\) steps of the timed interpreter reach the borrow phase with the borrow flag set and the clock head at \(0\); the retained action bits and halting flag are unchanged.
From the borrow phase with a given carry flag and the clock head at \(0\), exactly as many steps of the timed interpreter as the clock word’s length replace the word by its borrow subtraction, ending at the right blank with the resulting underflow flag in the control; the word’s width is preserved.
From the six-tape lift of a pending-action configuration with clock word \(w\), exactly \(2|w| + 2\) steps of the timed interpreter complete the clock rewind and the borrow pass, reaching the borrow phase at the right blank with the subtracted word and its underflow flag in the control.
From the six-tape lift of a four-tape configuration whose control is a pending-action state, one step of the timed interpreter enters the clock-rewind phase — retaining the action’s decoded bits and halting flag — with the clock head moved to the last clock bit, and without applying any source transition.
If the clock word \(w\) has positive little-endian value, then from the six-tape lift of a pending-action configuration, exactly \(2|w| + 4\) steps of the timed interpreter decrement the budget once and apply the retained action: the result is the lift of one step of the unstopped table interpreter, with clock word the borrow subtraction of \(w\).
If the clock word \(w\) has little-endian value \(0\), then after \(2|w| + 3\) steps of the timed interpreter from the six-tape lift of a pending-action configuration, the machine has halted with real output exactly the single timeout tag \(0\), regardless of any emissions buffered so far.
For every effective coding scheme, code string \(\alpha \), input \(x\), and deadline \(t\), the timed universal machine on the nested pair encoding of the binary representation of \(t\), the code \(\alpha \), and the input \(x\) halts with output the deadline-inclusive answer of the decoded machine’s initial configuration at budget \(t\), within \((\text{startup bound} + \text{block bound} + 14)\, (t+1)^2\) steps, where both bounds depend only on \(\alpha \).
From the six-tape lift of a pending-action configuration, with the control replaced by the execute phase, one step of the timed interpreter applies the retained source action: the result equals the lift of one step of the unstopped table interpreter from the four-tape configuration, with the clock word unchanged.
On the nested pair encoding \(\langle \langle b, \alpha \rangle , x\rangle \), the timed universal machine reaches, within \(4|b|\) plus the code-only startup bound steps, a framed configuration over some inactive tapes and heads: the six-tape lift, with clock word \(b\), of the interpreter checkpoint of the decoded machine’s initial configuration on \(x\), with the table cursor just past the serialization header (at twice the binary state-count width, plus \(2\), plus the initial state index, plus \(1\)).
For all bit strings \(b\), \(\alpha \), \(x\), the nested pair encoding \(\langle \langle b, \alpha \rangle , x\rangle \) equals the physical clock prefix of \(b\) (each bit quadrupled, then the four cells \(0011\)) followed by the pair encoding \(\langle \alpha , x\rangle \).
Let \(M\) be a coded machine with serialization of length \(L\) and state count \(N+1\), let \(\mathrm{src}\) be any configuration of \(M\), let the table cursor be at most \(L\), and let the clock word have width \(w\) and little-endian value \(r\). Then the timed interpreter, started from the six-tape lift of the interpreter checkpoint of \(\mathrm{src}\) with that clock word, halts within
steps, and its real output is exactly the deadline-inclusive answer of \(\mathrm{src}\) at budget \(r\): halting on the final allowed transition counts as success.
Suppose a machine’s transitions, on a family of control states indexed by a register type, always perform a right-moving single-lane write determined by a local visit function of the register and the symbol read on the designated lane. Then running the machine for exactly the length of a word laid out on that lane’s zipper computes the left fold of the visit function over the word: the final control carries the folded register, the frontier advances by the word’s length, the processed stack holds the rewritten word reversed, and all inactive tapes are unchanged.
A configuration in which one designated work tape holds a finite zipper — a frontier position, a stack of processed cells, and a list of unprocessed cells — while every other component (the remaining tapes, all other heads, the input position, and the output) is taken from a base configuration.
Applying a right-moving single-lane action that writes a symbol \(b\) to a single-active-lane configuration whose unprocessed list begins with a cell \(a\) consumes \(a\), pushes \(b\) onto the processed stack, advances the frontier by one, and passes to the designated control; all inactive data is unchanged.
For every \(j\) at most the buffered output’s length, from the buffer-rewind phase with the buffer head at \(j-1\), exactly \(j+1\) steps of the timed interpreter reach the flush phase with the buffer head at \(0\), emitting the success tag \(1\) at the left blank before any data bit.
If the buffered output splits as a processed prefix \(l\) followed by a remainder \(r\), with the buffer head at \(|l|\), then the flush phase of the timed interpreter emits exactly the bits of \(r\), in order, in \(|r| + 1\) steps and halts at the right blank without a further emission.
Suppose that, whenever the output-buffer lane reads its designated symbol, the timed interpreter’s transition in control \(q\) is a pure buffer move by \(d \in \{ -1,0,1\} \) with an optional real emission. Then one step from the output-phase configuration moves only the buffer head by \(d\), appends the optional emission to the real output, and passes to the designated control.
If the stopped table interpreter has not halted after \(t\) steps from a four-tape configuration, then \(t\) steps of the timed interpreter from the six-tape lift of that configuration equal the lift of the \(t\)-step stopped run, with an unchanged clock word; no countdown or output phase is visited in the interior of the run.
For every effective machine-coding scheme there is a single finite-tape machine \(U\) over \(\{ 0,1\} \) such that for every code string \(\alpha \) there is a constant \(C\) with the following two properties for every input \(x\) and every deadline \(t \in \mathbb {N}\). If the machine denoted by \(\alpha \) halts on \(x\) within \(t\) steps with output \(w\), then \(U\) on the nested pair encoding of the binary representation of \(t\), the code \(\alpha \), and the input \(x\) halts with output \(1\) followed by \(w\) within \(C (t+1)^2\) steps. If instead the denoted machine produces no output on \(x\) within \(t\) steps, then \(U\) halts with output the single timeout tag \(0\) within the same bound. Halting is checked after every simulated transition including the \(t\)-th, so halting exactly at the deadline counts as success.
Let \(c\) be an effective machine-coding scheme whose canonizer runs in time \(T \colon \mathbb {N} \to \mathbb {N}\); for a bit string \(\alpha \) let \(M_\alpha \) be the machine \(\alpha \) decodes to, \(L_\alpha \) the length of its canonical serialization, \(N_\alpha \) its number of states and \(q_\alpha \) the index of its initial state, write \(\mathrm{bin}(n)\) for the binary representation of \(n\) (least significant bit first) and \(\langle a, b\rangle \) for the pair encoding of bit strings (each bit of \(a\) doubled, then \(01\), then \(b\)), and set
Then there is a single finite-tape machine \(U\) over \(\{ 0,1\} \) such that for every bit string \(\alpha \), input \(x\) and deadline \(t \in \mathbb {N}\): if \(M_\alpha \) halts on \(x\) within \(t\) steps with output \(w\), then \(U\) on \(\langle \langle \mathrm{bin}(t), \alpha \rangle , x\rangle \) halts within \(C_\alpha (t+1)^2\) steps with output \(1\) followed by \(w\); and if there is no string \(w\) such that \(M_\alpha \) halts on \(x\) within \(t\) steps with output \(w\), then \(U\) on \(\langle \langle \mathrm{bin}(t), \alpha \rangle , x\rangle \) halts within \(C_\alpha (t+1)^2\) steps with output the single bit \(0\).
The lift of a four-tape interpreter action to six tapes: the four core lanes act as before, the clock lane is untouched, any emission is written to the output-buffer lane (advancing its head) instead of being emitted, and a halting successor is redirected to the emission-start phase so that the machine stays live.
Applying the six-tape lift of a four-tape interpreter action to the six-tape lift of a four-tape configuration (with any clock word) equals the lift of the configuration obtained by applying the action: the clock word is unchanged and any emission is captured on the buffer lane, including an emission on the action’s halting transition.
A six-tape action that leaves the four core lanes and the input head untouched, performs a designated optional write and move on the clock lane and on the output-buffer lane, optionally emits one real output bit, and passes to a designated control.
Fixed-width binary subtraction of an incoming borrow flag from a little-endian word: each output bit is the exclusive or of the corresponding input bit with the propagated borrow, the borrow propagates exactly through the zero bits, and the result is the final (underflow) borrow flag together with a word of the same width.
After \(4|b| + 3|\alpha | + 8\) plus the canonizer’s time bound at \(|\alpha |\) steps on the nested pair encoding, the composed clock-preserving canonizer has halted; its output is the serialization of the machine decoded from \(\alpha \), its input head is parked at position \(4|b| + 2|\alpha | + 7\), and its retained clock lane still holds the buffer of \(b\) with head at \(|b|\).
For every \(t\), running the composed clock-preserving canonizer \(4|b| + 3|\alpha | + 8 + t\) steps on the nested pair encoding equals, for some internal rewind flag, the buffered second-phase configuration of the canonizer run for \(t\) steps on virtual input \(\alpha \); the physical input head and the retained clock word \(b\) remain stationary.
On the nested pair encoding \(\langle \langle b, \alpha \rangle , x\rangle \), exactly \(4|b| + 3|\alpha | + 8\) steps of the composed clock-preserving canonizer complete the parse and hand the extracted code \(\alpha \) to the canonizer in its initial configuration, with the clock word \(b\) preserved on its lane (head at \(|b|\)) and the physical input head parked at the start of the suffix region.
If the base machine halts on input \(x\) within \(T\) steps, then for some \(t \le T + 1\), running the capture wrapper \(t\) steps from its initial configuration on \(x\) reaches the interpreter-entry configuration for the base machine’s configuration at time \(T\); since halting is absorbing, all fields of that configuration — including the parked input head — are retained, not merely the completed output.
For every live configuration of the base machine, one step of the capture wrapper from its first-phase embedding equals the embedding of one step of the base machine: the input head and the original work tapes move in lockstep, any emission — including one on the base machine’s halting transition — is appended to the captured table rather than the real output, and the later fresh lanes stay untouched; administrative states remain live.
The first-phase configuration of the capture wrapper corresponding to a configuration of the base machine \(M\): \(M\)’s tapes and heads are embedded, the captured-table lane holds \(M\)’s accumulated output as a buffer with head at its right blank, the remaining fresh lanes are blank, and the real output is empty.
Let a halted configuration of the base machine hold, on the designated clock tape, the buffer of a clock word with head at the word’s length. Then its interpreter-entry capture configuration equals the frame — over the base configuration’s tapes and heads — of the six-tape lift, with that clock word, of the stopped interpreter’s entry configuration on the captured table (the base configuration’s accumulated output) with the same parked input head.
The interpreter-entry configuration of the capture wrapper for a halted configuration of the base machine: identical to its first-phase embedding except that the finite control is the timed interpreter’s initial state; the retained clock lane becomes active again during interpretation.
The capture wrapper of a machine \(M\) with a designated clock tape: it runs \(M\) in lockstep while writing \(M\)’s emissions to a fresh table lane instead of the output; when \(M\) halts, control transfers to the timed interpreter, which runs on the captured table, the retained clock lane, and four further fresh lanes.
Suppose that, whenever the clock lane reads its designated symbol, the timed interpreter’s transition in control \(q\) is a pure clock move by \(d \in \{ -1,0,1\} \) (no writes, no emission) into a control \(q'\). Then one step from the clock-phase configuration moves only the clock head by \(d\) and enters \(q'\).
The finite control of the timed interpreter: a working copy of the table interpreter’s control, clock phases (rewind and borrow) that retain a selected action’s eight decoded bits and halting flag while the countdown is serviced, an execute phase, and three output phases (emission start, buffer rewind, and flush).
For a coded one-work-tape binary machine and a state \(q\), the list of the nine transition actions of \(q\), one for each pair of an optional input read and an optional work read, ordered with the input read (blank, \(0\), \(1\)) as the major key and the work read (in the same order) as the minor key.
For every coded machine, state \(q\), and pair of optional input and work reads, the nine-action list of \(q\) has length \(9\), and its entry at the record index determined by the two reads (a number in \(\{ 0,\dots ,8\} \)) is exactly the machine’s transition action for \(q\) on those reads.
Applying an administrative interpreter action — one that moves the table cursor by \(d_t \in \{ -1,0,1\} \), optionally writes a symbol at the state-tape cursor, moves the state cursor by \(d_s\), and passes to a control state \(q'\) — to an interpreter evaluation configuration changes exactly the two designated cursors and, when a write is present, the state-tape cell at the cursor; the input head, the mirrored data, and the accumulated output are unchanged.
Suppose the table holds, starting at the cursor position \(|l|\), the doubling \(b_1 b_1 b_2 b_2 \cdots \) of a count string of \(n\) bits followed by the delimiter pair \(01\). Then the count-parsing phase of the stopped table interpreter traverses this field in exactly \(2n+2\) steps, advancing the cursor past the delimiter and entering the initial-state copier (when copying was requested) or the initial-state skipper; no binary arithmetic on the count’s value is performed.
Consider an interpreter evaluation configuration, whose four work tapes are: a table tape holding the buffer of a bit string \(T\) with cursor \(p\), a state tape \(\sigma \colon \mathbb {Z} \to \{ 0,1,\text{blank}\} \) with cursor \(s\), and a mirrored work tape and a boundary-marker tape taken from a base configuration. The four symbols under its work heads are, in order, the table symbol at \(p\), the state symbol \(\sigma (s)\), and the two symbols read on the corresponding tapes of the base configuration.
If, on the symbols read in an interpreter evaluation configuration, the transition function of the stopped table interpreter returns the administrative action with table move \(d_t\), optional state-tape write, state move \(d_s\), and next control \(q'\), then one machine step from that configuration yields exactly the evaluation configuration with control \(q'\), table cursor shifted by \(d_t\), and state tape and state cursor updated accordingly.
If the table holds \(n\) ones followed by a zero starting at the cursor position \(|l|\), then the initial-state-skipping phase of the stopped table interpreter traverses this unary field in exactly \(n+1\) steps and enters record-group selection with the table cursor at \(|l|+n+1\).
The first step of the stopped table interpreter from its entry configuration installs the two permanent markers (the state-tape origin marker and the virtual-input boundary marker) and enters the unconditional table rewind with the table cursor on the last table cell.
Suppose the captured table \(T\) is the pair encoding of a count string of \(m\) bits with a body consisting of \(n\) ones, a terminator zero, and a record region. Then exactly \(|T| + 2m + 2n + 7\) steps of the stopped table interpreter bring the entry configuration to the main simulation phase, with the table cursor at \(2m + 2 + n + 1\) (just past the header) and the state tape holding the unary representation of \(n\) with cursor \(1\); every intermediate configuration keeps the input head fixed and the output empty, and no transition-record lookup is involved.
The inactive part of an interpreter configuration during initialization: the input head parked at a given position, a blank mirrored work tape, the virtual-input left-boundary marker installed at cell \(0\) of its tape with its head at \(1\) (also correct for an empty input suffix), and empty output.
The configuration of the stopped table interpreter at entry: control in the start phase, the captured table held on the first work tape as a buffer with its cursor at the right blank, three blank auxiliary tapes, empty output, and the input head already parked at a given position of the input.
Let \(M\) be a coded machine with serialization of length \(L\) and state count \(N+1\), let \(\mathrm{src}\) be a non-halted configuration of \(M\), and let the table cursor \(p\) satisfy \(p \le L\). Then there exist a duration \(d \le 3L + 5(N+1) + 20\), a new cursor \(p' \le L\), and a configuration \(\mathrm{ready}\) whose control is a pending-action state, such that the stopped table interpreter runs from the interpreter checkpoint of \(\mathrm{src}\) to \(\mathrm{ready}\) in exactly \(d\) steps, and one step of the unstopped interpreter from \(\mathrm{ready}\) is the interpreter checkpoint of the configuration obtained by one step of \(M\) from \(\mathrm{src}\), with table cursor \(p'\).
For every coded machine, state \(q\), and pair of optional input and work reads, the serialization splits as: the header, then the record groups of the \(q\) states preceding \(q\) (each a list of nine records), then the records of \(q\)’s group preceding the read offset selected by the two reads, then the serialized record of the machine’s transition action for \(q\) on those reads, then a remainder.
Suppose the table holds, at the cursor, the unary successor field of a record (its number of ones encoding the optional successor state) followed by its terminator, and the state tape is the unary tape for \(0\) with cursor \(1\). Then within exactly the successor-preparation cost, the stopped table interpreter installs the unary representation of the successor index (or of \(0\) in the halting case) on the state tape with cursor \(1\) and enters the pending-action phase carrying the eight decoded bits and the halting flag, with the table cursor stopped on the field’s terminator.
Suppose the table holds a word \(b\) of eight bits starting at position \(|l|\), the action-reading phase of the stopped table interpreter is at field index \(i\) with \(n = 7-i\) fields left, and the accumulated eight-bit register agrees with \(b\) on all indices below \(i\). Then exactly \(n+1\) steps fill the register with all of \(b\) and enter the successor-decoding phase, with the table cursor at \(|l|+8\).
Suppose the table holds, at the cursor, the concatenated records of \(g\) preceding state groups (nine records each) followed by the records preceding the selected read offset, and the state tape is the unary tape for \(g\) with cursor \(1\). Then in exactly \(2g\) plus the total serialized length of the traversed records plus \(3\) steps, the stopped table interpreter destructively counts down the unary state while skipping its groups, rewinds the erased state tape, skips the read-offset prefix (entering the action reader directly when the offset is zero), and arrives in the action-reading phase with a cleared register, the table cursor at the selected record, and the state tape reset to the unary tape for \(0\).
In the fixed-field-skipping phase of the stopped table interpreter, with current field index \(i\) and \(n\) further fields remaining (so \(i + n = 7\)), exactly \(n+1\) steps advance the table cursor by \(n+1\) cells and enter the unary-tail scanner; no tape content is inspected or modified.
Suppose the table holds, at the cursor, the concatenated records of a list of \(g\) state groups, each consisting of nine records, and the state tape holds a partially erased unary state whose remaining ones number exactly \(g\). Then in exactly \(g\) plus the total serialized length plus one steps, the stopped table interpreter erases one unary state symbol per group, skips all of the groups’ records, and enters the state rewind, with the table cursor past the traversed records and all erasures recorded on the state tape.
If a complete serialized transition record — eight fixed bits followed by a unary successor field and its terminator — occupies the table starting at the cursor, then the skipping phase of the stopped table interpreter traverses it in exactly its serialized length, decrementing the bounded record counter, or completing the skip request when the counter is zero, with the cursor moved past the record.
If the table holds, at the cursor, the concatenated serializations of a nonempty list of transition records whose length is one more than the bounded record counter, then the skipping phase of the stopped table interpreter traverses all of them in exactly the total serialized length and finishes in the requested continuation with the cursor past the last record; no extra transition occurs between consecutive records.
If the table holds \(n\) ones followed by a zero at cursor position \(|l|\), the unary-skipping phase of the stopped table interpreter traverses them in exactly \(n+1\) steps, ending with the cursor at \(|l|+n+1\); on the terminator it completes the skip request when the bounded record counter is zero, and otherwise decrements the counter and re-enters fixed-field skipping.
Suppose the state tape holds the marker symbol \(0\) at cell \(0\) and no cell at a positive position holds \(0\), and suppose the current control of the stopped table interpreter moves the state head left on every non-marker read while, on reading the marker, it moves right once and transfers to a designated continuation control \(q'\). Then from state cursor \(j \in \mathbb {N}\) the interpreter reaches \(q'\) with state cursor \(1\) in exactly \(j+1\) steps, without changing any tape contents, the input head, or the output; the premise is independent of whether the traversed cells are erased blanks or retained unary symbols.
Suppose the stopped table interpreter is in its table-rewind phase with the table cursor at \(j-1\), where \(j\) is at most the table length. Then exactly \(j+1\) steps bring it to the count-parsing phase with the table cursor at \(0\), leaving the state tape, the mirrored data, the input head, and the accumulated output unchanged.
Suppose the current control \(q\) of the stopped table interpreter appends one unary symbol to the state tape and advances both cursors whenever the table reads \(1\), and on reading \(0\) transfers to a continuation control \(q'\) with a designated final table move. If the table holds \(n\) ones followed by a zero at cursor position \(|l|\), then \(n+1\) steps transform the unary state tape for \(j\) (state cursor \(j+1\)) into the unary state tape for \(j+n\) (state cursor \(j+n+1\)), with the table cursor ending at \(|l|+n\) plus the final move. This single gadget serves both initial-state extraction and successor-state replacement.
The four-tape table interpreter of the universal-machine construction, modified so that every control state holding a fully decoded transition record halts immediately instead of applying the record; on all other control states its transition function agrees with the original interpreter.
The configuration of the assembled capture machine determined by a six-tape timed-interpreter configuration together with inactive tapes and heads for the old block: the interpreter’s clock lane is placed at the retained clock index of the old block, its remaining five lanes occupy the fresh indices, and every other old lane carries the given inactive data.
The lift of a six-tape timed-interpreter action to the assembled machine with \(k + 5\) tapes, for a base machine with \(k\) tapes and a designated clock tape: the action’s clock component acts on the retained clock lane of the old block, its remaining five components act on the five fresh lanes, and all other old lanes are left untouched.
The six-tape interpreter with a countdown clock: it follows the table interpreter, buffering all emissions, until a decoded action is pending; it then rewinds the clock lane and performs one fixed-width borrow pass. A successful borrow executes the retained action; an underflow halts the machine emitting only the timeout tag \(0\). When the simulated machine halts, the machine emits the success tag \(1\) and flushes the buffered output before halting.
The six-tape configuration representing a four-tape interpreter configuration together with a clock word: the four core lanes are copied, the clock lane holds the clock word as a buffer with head at its right blank, the buffer lane holds the four-tape configuration’s accumulated output, and the real output is empty; a halted four-tape control maps to the emission-start phase.
Suppose the prefix parser’s transition in control \(q\), on the input symbol found at physical position \(p\) of the nested pair encoding, moves the input head right, writes nothing to the work tape, optionally emits one bit, and passes to a control \(q'\). Then one machine step from the corresponding parser configuration advances the physical input position by one, appends the optional emission to the output, keeps the stored clock unchanged, and enters \(q'\).
For every \(j \le |b|\), running the prefix parser \(4j\) steps from its initial configuration on the nested pair encoding \(\langle \langle b, \alpha \rangle , x\rangle \) stores exactly the first \(j\) clock bits on its work tape, leaves the output empty, and places the input head just past the \(4j\) consumed cells, still in the first clock phase.
Running the prefix parser \(4|b| + 4\) steps from its initial configuration on the nested pair encoding \(\langle \langle b, \alpha \rangle , x\rangle \) consumes the entire quadrupled clock and its four separator cells, stores exactly \(b\) on the work tape, produces no output, and enters the code-extraction phase.
For every \(j \le |\alpha |\), running the prefix parser \(4|b| + 4 + 2j\) steps from its initial configuration on the nested pair encoding \(\langle \langle b, \alpha \rangle , x\rangle \) keeps the stored clock \(b\) intact, emits exactly the first \(j\) code bits as output, and remains in the code-extraction phase with the input head just past the consumed cells.
Running the prefix parser \(4|b| + 2|\alpha | + 6\) steps from its initial configuration on the nested pair encoding \(\langle \langle b, \alpha \rangle , x\rangle \) halts it with the full clock \(b\) stored on the work tape, output exactly \(\alpha \), and the input head parked at the start of the suffix region holding \(x\); both delimiters are consumed, also when \(b\) or \(\alpha \) is empty.
The configuration of the prefix parser on the nested pair encoding \(\langle \langle b, \alpha \rangle , x\rangle \) with a given control state and input position, the captured clock string held on its work tape as a buffer with head at the right blank, and a given accumulated output.
A one-work-tape machine over \(\{ 0,1\} \) that scans the nested pair encoding from the left: it stores each undoubled clock bit on its work tape (one write per four input cells), then emits each undoubled code bit as output (one emission per two input cells), and halts on the outer separator, before reading any of the input suffix.
The part of the timed universal machine’s startup cost that depends only on the code string \(\alpha \): three times \(|\alpha |\), plus the canonizer’s time bound at \(|\alpha |\), plus the length of the decoded machine’s serialization, plus twice the width of its binary state count, plus twice its initial state index, plus \(16\).
For a natural number \(m\), the unary tape of \(m\) is the tape \(\mathrm{ones}_m : \mathbb {Z} \to \{ 0, 1, \sqcup \} \), where \(\sqcup \) denotes the blank symbol, given by
Let \(m \in \mathbb {N}\), and let \(\mathrm{ones}_m : \mathbb {Z} \to \{ 0, 1, \sqcup \} \) be its unary tape, which holds \(1\) on cells \(0, \dots , m-1\) and the blank \(\sqcup \) elsewhere. If \(c \in \mathbb {Z}\) satisfies \(m \le c\), then \(\mathrm{ones}_m(c) = \sqcup \).
Let \(m \in \mathbb {N}\), and let \(\mathrm{ones}_m : \mathbb {Z} \to \{ 0, 1, \sqcup \} \) be its unary tape, which holds \(1\) on cells \(0, \dots , m-1\) and the blank \(\sqcup \) elsewhere. If \(c \in \mathbb {Z}\) satisfies \(0 \le c\) and \(c \lt m\), then \(\mathrm{ones}_m(c) = 1\).
Let \(m \in \mathbb {N}\), and let \(\mathrm{ones}_m : \mathbb {Z} \to \{ 0, 1, \sqcup \} \) be its unary tape, which holds \(1\) on cells \(0, \dots , m-1\) and the blank \(\sqcup \) elsewhere. If \(c \in \mathbb {Z}\) satisfies \(c \lt 0\), then \(\mathrm{ones}_m(c) = \sqcup \).
For \(m \in \mathbb {N}\), let \(\mathrm{ones}_m : \mathbb {Z} \to \{ 0, 1, \sqcup \} \) be the unary tape of \(m\), which holds \(1\) on cells \(0, \dots , m-1\) and the blank \(\sqcup \) elsewhere. The unary tape of \(0\) is the all-blank tape: \(\mathrm{ones}_0(c) = \sqcup \) for every \(c \in \mathbb {Z}\).
For \(m \in \mathbb {N}\), let \(\mathrm{ones}_m : \mathbb {Z} \to \{ 0, 1, \sqcup \} \) be the unary tape of \(m\), which holds \(1\) on cells \(0, \dots , m-1\) and the blank \(\sqcup \) elsewhere. For every \(j \in \mathbb {N}\), the tape obtained from \(\mathrm{ones}_{j+1}\) by setting cell \(j\) to the blank \(\sqcup \) and leaving every other cell unchanged is \(\mathrm{ones}_j\).
For \(m \in \mathbb {N}\), let \(\mathrm{ones}_m : \mathbb {Z} \to \{ 0, 1, \sqcup \} \) be the unary tape of \(m\), which holds \(1\) on cells \(0, \dots , m-1\) and the blank \(\sqcup \) elsewhere. For every \(m \in \mathbb {N}\), the tape obtained from \(\mathrm{ones}_m\) by setting cell \(m\) to \(1\) and leaving every other cell unchanged is \(\mathrm{ones}_{m+1}\).
Let \(T : \mathbb {Z} \to \{ 0, 1, \sqcup \} \) be a tape, where \(\sqcup \) denotes the blank symbol, let \(p \in \mathbb {Z}\) be a head position, and let \(w\) be an optional write, which is either “no write” or a symbol \(s \in \{ 0, 1, \sqcup \} \) to be written. The tape after the write is \(T\) itself if \(w\) is “no write”. If \(w\) writes \(s\), it is the tape that holds \(s\) at cell \(p\) and agrees with \(T\) at every other cell.
For a Boolean word \(v\), \(\mathrm{tok}(v)\) is the pair (token, remainder) obtained by splitting off the maximal leading block of \(\mathrm{true}\)’s together with the \(\mathrm{false}\) that terminates it: if \(v = \mathrm{true}^m \, \mathrm{false} \, r\) with \(m \ge 0\), the result is \((\mathrm{true}^m \, \mathrm{false}, \, r)\), and if \(v = \mathrm{true}^m\) contains no \(\mathrm{false}\) (including \(v\) empty), the result is \((\mathrm{true}^m, \varepsilon )\) with \(\varepsilon \) the empty word.
For every effective machine-coding scheme there is a single finite-tape machine \(U\) over the alphabet \(\{ 0,1\} \) such that for every code string \(\alpha \) there is a constant \(C\) (depending on \(\alpha \)) with the following two properties for every input string \(x\). Forward: whenever the coded machine denoted by \(\alpha \) halts on \(x\) within \(t\) steps with output \(w\), the machine \(U\) on the pair encoding \(\langle \alpha , x\rangle \) (each bit of \(\alpha \) doubled, the delimiter \(01\), then \(x\) verbatim) halts with the same output \(w\) within \(C \cdot (t+1)\) steps. Converse: every output that \(U\) produces on \(\langle \alpha , x\rangle \) upon halting is an output the denoted machine produces on \(x\) upon halting, so divergence is preserved.
Let \(M\) be a coded machine, \(\alpha \) a code string, \(s\) a configuration of \(M\)’s one-work-tape form on an input \(x\), and \(a\) a transition action of \(M\). One step of the table interpreter from the evaluation configuration over the interpreter checkpoint of \(s\) (at an arbitrary table cursor) whose control is the record-application phase carrying the eight fixed bits of \(a\) and the flag recording whether \(a\) halts, whose table is the canonical serialization of \(M\) at a cursor \(p\), and whose state tape is the intact unary encoding of \(a\)’s successor index (\(0\) when \(a\) halts) with state cursor \(1\), is exactly the interpreter checkpoint of the configuration obtained by applying \(a\) to \(s\), with table cursor \(p\).
Let \(M\) and \(U\) be multi-tape machines over \(\{ 0,1\} \) on inputs \(x\) and \(y\) respectively, let \(R\) be a relation between their configurations, and let \(S, B \in \mathbb {N}\). Assume (startup) there is a \(t \le S\) such that the configuration of \(U\) after \(t\) steps from its initial configuration is related to \(M\)’s initial configuration, and (step) whenever a source configuration is related to a target configuration, there is a duration \(d\) with \(1 \le d \le B\) such that running \(U\) for \(d\) further steps yields a configuration related to the source’s successor. Then for every \(n\) there is a time \(t\) with \(n \le t \le S + Bn\) such that the configuration of \(U\) after \(t\) steps is related to the configuration of \(M\) after \(n\) steps. The lower bound \(n \le t\) makes the run cofinal, without assuming that either machine eventually halts.
Let \(c\) be an effective machine-coding scheme with canonizer time bound \(T_c\), let \(D\) be a four-tape interpreter, and let \(\alpha \) and \(x\) be bit strings. Start the output-capture wrapper of the composed prefix-extraction/canonizer machine of \(c\), with interpreter \(D\), on the self-delimiting pairing \(\langle \alpha , x\rangle \). Then there is a \(t \le 3\left\lvert \alpha \right\rvert +5+T_c(\left\lvert \alpha \right\rvert )\) such that after \(t\) transitions the control is the interpreter’s initial state, the physical input head is at position \(2\left\lvert \alpha \right\rvert +3\) (the start of the verbatim suffix), the interpreter’s first work tape holds — written from coordinate \(0\) and blank elsewhere — the serialization of the canonical transition table of the machine that \(c\) decodes from \(\alpha \), and the real output is still empty.
Suppose the table is \(l \cdot d(\beta ) \cdot 01 \cdot r\), where \(d(\beta )\) doubles each bit of a bit string \(\beta \) and \(01\) is the unequal delimiter pair. Starting the interpreter’s count parser (carrying a Boolean flag and a record index) at table cursor \(|l|\), exactly \(2|\beta | + 2\) steps move the table cursor past the doubled count field and the delimiter to position \(|l| + 2|\beta | + 2\), entering the initial-state copier when the flag is set and the initial-state skipper with the carried index otherwise. No binary arithmetic on the value of \(\beta \) is performed, and the state tape, state cursor, and all other data are unchanged.
Let \(c\) be an effective machine-coding scheme, \(U\) a finite-tape machine over \(\{ 0,1\} \), and \(S, B\) functions assigning a bound to each code string. For each code string \(\alpha \) and input \(x\), let a relation be given between configurations of the machine decoded from \(\alpha \) (in one-work-tape form) on \(x\) and configurations of \(U\) on the pair encoding \(\langle \alpha , x \rangle \). Assume: (startup) for all \(\alpha , x\), within some \(t \le S(\alpha )\) steps \(U\) reaches a configuration related to the decoded machine’s initial configuration; (step) every related pair advances — one source step against some \(d\) target steps with \(1 \le d \le B(\alpha )\) — to a related pair; (halting) related configurations are halted simultaneously; and (output) related configurations have equal accumulated output. Then for every \(\alpha \) there is a constant \(C\) such that for every input \(x\): whenever the decoded machine computes an output within \(t\) steps on \(x\), the machine \(U\) computes the same output on \(\langle \alpha , x \rangle \) within \(C(t+1)\) steps; and conversely, every output \(U\) computes on \(\langle \alpha , x \rangle \) within some time is computed by the decoded machine on \(x\) within some time.
For every coded machine \(M\) with state count \(m\) and initial state index \(q_0\): writing \(\ell \) for the number of binary digits of \(m\) and \(L\) for the length of the canonical serialization of \(M\), one has \(2\ell + 2 + q_0 + 1 \le L\); that is, the position just past the doubled count field, its delimiter pair, and the unary initial-state field with its terminator lies within the serialized table.
Suppose the table is \(l \cdot 1^n \cdot 0 \cdot r\). From the interpreter’s initial-state skipping phase (carrying a record index) with table cursor \(|l|\), exactly \(n + 1\) steps move the table cursor past the unary field and its terminating \(0\) to position \(|l| + n + 1\) and enter the record-group phase with the carried index; the state tape, state cursor, and all other data are unchanged.
Let \(c\) be an effective machine-coding scheme, \(\alpha \) a code string, and \(x\) an input string, and let \(N\) denote the machine decoded from \(\alpha \), with \(\ell \) the number of binary digits of its state count and \(q_0\) its initial state index. Then there exist a time \(t\) not exceeding the startup bound for \(\alpha \) (which is independent of \(x\)) and contents and head positions for the inactive canonizer tapes such that running the complete universal-machine candidate for \(t\) steps from its initial configuration on the pair encoding \(\langle \alpha , x \rangle \) yields exactly the right-block embedding, over those inactive tapes, of the interpreter checkpoint representing the initial configuration of \(N\) on \(x\) with table cursor \(2\ell + 2 + q_0 + 1\) — that is, with the canonical table captured, its cursor parked just past the header, the initial state in unary, a blank simulated work tape, and the virtual-input boundary marker installed.
Let \(M\) be a multi-tape Turing machine, \(c\) one of its configurations on some input, and \(s \le t\) natural numbers. If the configuration reached from \(c\) after \(t\) steps is not halted, then neither is the configuration reached after \(s\) steps: since halting is absorbing, a live later configuration excludes all earlier halts.
Let \(M\) be a coded machine with state count \(N\) and canonical serialization of length \(L\), let \(\alpha \) be a code string, let \(s\) be a configuration of \(M\)’s one-work-tape form on an input \(x\) that has not halted, and let \(p \le L\) be a table cursor. Then there exist a duration \(d\) with \(1 \le d \le 3L + 5(N + 1) + 20\) and a cursor \(p' \le L\) such that running the table interpreter for \(d\) steps from the interpreter checkpoint of \(s\) with table cursor \(p\) yields exactly the interpreter checkpoint of the configuration obtained from \(s\) by one step of \(M\), with table cursor \(p'\).
For every coded machine \(M\), state \(q\) of \(M\), and pair of reads (an input symbol and a work symbol, each blank, \(0\), or \(1\)), the canonical serialization of \(M\) decomposes as: the header, then the concatenated serialized records of a list of exactly \(q\) groups of nine actions each (the full groups of the states preceding \(q\)), then the concatenated serialized records of a list of actions whose length is the record index of the read pair, then the serialized record of the action \(M\) performs at \(q\) on those reads, then a remainder.
Consider the tape that carries the single-cell word \(1\) at coordinate \(0\) and is blank at every other coordinate, including all negative ones. For every natural number \(p\), the cell of this tape at coordinate \(p\) contains the mark \(1\) if \(p = 0\) and is blank otherwise.
Let \(\alpha \) and \(x\) be bit strings and let \(j \lt \left\lvert \alpha \right\rvert \). In the self-delimiting pairing \(\langle \alpha , x\rangle \) (each bit of \(\alpha \) doubled, then the separator pair \(01\), then \(x\) verbatim), the cells of index \(2j\) and \(2j+1\) (zero-indexed) both contain the \(j\)-th bit of \(\alpha \); both cells of a doubled code bit therefore lie before the separator.
For all bit strings \(\alpha \) and \(x\), the self-delimiting pairing \(\langle \alpha , x\rangle \) — each bit of \(\alpha \) written twice, followed by the separator pair \(01\), followed by \(x\) verbatim — has length exactly \(2\left\lvert \alpha \right\rvert + 2 + \left\lvert x\right\rvert \). In particular the extent of the code region is independent of the suffix \(x\).
For all bit strings \(\alpha \) and \(x\), the cells of index \(2\left\lvert \alpha \right\rvert \) and \(2\left\lvert \alpha \right\rvert +1\) in the self-delimiting pairing \(\langle \alpha , x\rangle \) (each bit of \(\alpha \) doubled, then the separator pair \(01\), then \(x\) verbatim) contain the bits \(0\) and \(1\) respectively: the separator begins immediately after the doubled code region.
For all bit strings \(\alpha \) and \(x\) and every natural number \(j\), the optional cell lookup at index \(2\left\lvert \alpha \right\rvert +2+j\) in the self-delimiting pairing \(\langle \alpha , x\rangle \) agrees with the optional lookup at index \(j\) in \(x\); both are simultaneously undefined when \(j \ge \left\lvert x\right\rvert \).
Suppose the table is \(l \cdot 1^u \cdot 0 \cdot r\), where \(u\) is the unary size of an optional successor state (\(0\) when the successor is absent and \(q + 1\) when it has index \(q\)). From the interpreter’s successor-decoding phase carrying eight buffered action bits, with table cursor \(|l|\), state tape the intact unary encoding of the index \(0\), and state cursor \(1\), exactly the successor-preparation cost (\(1\) when the successor is absent and \(2q + 4\) otherwise) in interpreter steps reaches the record-application phase carrying the same bits and a flag recording whether the successor is absent, with table cursor \(|l| + u\), state tape the intact unary encoding of the successor index (\(0\) when absent), and state cursor \(1\).
For every effective machine-coding scheme there is a single machine \(U\) over \(\{ 0,1\} \) with the following property: whenever a finite-tape binary machine computes a total function \(f\) on bit strings within a time bound \(T \colon \mathbb {N} \to \mathbb {N}\), there are a code string \(\alpha \) and a constant \(C\) such that for every input \(x\), the machine \(U\) on the pair encoding \(\langle \alpha , x\rangle \) halts with output \(f(x)\) within \(C \cdot (T(|x|)+1)^2\) steps.
Suppose the table is \(l \cdot b_0 b_1 \cdots b_7 \cdot r\) for bits \(b_0, \dots , b_7\). Consider the interpreter’s action-reading phase at a field \(f \in \{ 0, \dots , 7\} \), with table cursor \(|l| + f\) and a bit register agreeing with \(b\) at every index below \(f\). Then exactly \(8 - f\) interpreter steps reach the successor-decoding phase carrying the complete register \(b_0, \dots , b_7\), with table cursor \(|l| + 8\); the state tape, state cursor, and all other data are unchanged.
The serialized record of any one-work-tape machine action decomposes as its eight fixed action bits, followed by a block of \(1\)s whose length is the unary size of the optional successor field (zero when the action halts, one more than the successor index otherwise), followed by a single terminating \(0\).
For a multi-tape machine over \(\{ 0,1\} \) and configurations \(a\), \(b\), \(c\) on a common input: if running the machine for \(s\) steps from \(a\) yields \(b\), and running it for \(t\) steps from \(b\) yields \(c\), then running it for \(s + t\) steps from \(a\) yields \(c\).
Suppose the table is \(l \cdot G \cdot B \cdot r\), where \(G\) concatenates the serialized transition records of \(g\) groups of nine one-work-tape machine actions each and \(B\) concatenates the serialized records of a list of actions whose length is a read index \(i \in \{ 0, \dots , 8\} \), and suppose the state tape is the intact unary encoding of \(g\) with state cursor \(1\). From the interpreter’s record-group phase carrying \(i\), with table cursor \(|l|\), exactly \(2g + |G| + |B| + 3\) interpreter steps reach the action-reading phase at field \(0\) with a cleared bit register, table cursor \(|l| + |G| + |B|\), state tape the intact unary encoding of \(0\), and state cursor \(1\): the two table scans cost their serialized lengths, and the destructive state count with its rewind costs \(2g + 3\).
The canonical serialization of a coded machine \(M\) equals its header followed by the concatenation, over every state of \(M\) in increasing order, of the concatenated serialized transition records of that state’s nine-action list (its actions on all read pairs, in the fixed input-major serialization order).
For every coded machine \(M\) with initial state index \(q_0\) and state count \(m\), there is a bit string \(\rho \) (the transition-record table) such that the canonical serialization of \(M\) equals the pair encoding of the binary digit string of \(m\) with the string \(1^{q_0} \cdot 0 \cdot \rho \); that is, the serialization begins with exactly its doubled count field and delimiter, then the initial state in unary with its terminator, then the record table.
Consider the interpreter’s fixed-field skipping phase, carrying a destination and a count of remaining records, with field register \(f \in \{ 0,\dots ,7\} \) and table cursor \(p\). If \(f + n = 7\), then exactly \(n + 1\) interpreter steps advance the table cursor to \(p + n + 1\) and enter the unary-skipping phase with the same destination and remaining count; no tape cell is inspected or modified, and the state tape, state cursor, and all other data are unchanged.
Suppose the table is \(l \cdot G \cdot r\), where \(G\) concatenates the serialized transition records of \(g\) groups of one-work-tape machine actions, each group containing exactly nine records, and suppose the state tape is the state window with \(j\) erased and \(g\) remaining unary cells, with state cursor \(j + 1\). From the interpreter’s record-group phase carrying a read index \(i \in \{ 0, \dots , 8\} \), with table cursor \(|l|\), exactly \(g + |G| + 1\) interpreter steps reach the state-rewind phase remembering \(i\), with table cursor \(|l| + |G|\), state tape the state window with \(j + g\) erased and no remaining unary cells, and state cursor \(j + g + 1\): each remaining unary state symbol is erased at the cost of one step plus a nine-record skip, and the terminating blank costs one further step.
Suppose the table is \(l \cdot \rho \cdot r\), where \(\rho \) is the serialized transition record of a one-work-tape machine action over \(\{ 0,1\} \) (its eight fixed action bits followed by its self-delimiting unary successor field). From the interpreter’s fixed-field skipping phase, carrying a destination and a count of remaining records in \(\{ 0, \dots , 8\} \), with field register \(0\) and table cursor \(|l|\), exactly \(|\rho |\) interpreter steps advance the cursor to \(|l| + |\rho |\) and pass control to the continuation after a skip request for the destination when the remaining count is \(0\), and to the fixed-field skipping phase with the count decremented and a cleared field register otherwise; the state tape, state cursor, and all other data are unchanged.
Let \(a_1, \dots , a_{m+1}\) be one-work-tape machine actions over \(\{ 0,1\} \) whose serialized transition records appear consecutively in the table after a prefix \(l\), and consider the interpreter’s fixed-field skipping phase carrying a destination and the remaining-record count \(m \in \{ 0, \dots , 8\} \), with field register \(0\) and table cursor \(|l|\). Then exactly as many interpreter steps as the total length of the \(m + 1\) serialized records advance the table cursor past all of them and enter the continuation after a skip request for the destination; the state tape, state cursor, and all other data are unchanged.
Suppose the table is \(l \cdot 1^n \cdot 0 \cdot r\). From the table interpreter’s unary-skipping phase, carrying a destination and a count of remaining records in \(\{ 0, \dots , 8\} \), with table cursor \(|l|\), exactly \(n + 1\) interpreter steps advance the cursor to \(|l| + n + 1\) and pass control to the continuation after a skip request for the destination when the remaining count is \(0\), and to the fixed-field skipping phase with the count decremented and a cleared field register otherwise; the state tape, state cursor, and all other data are unchanged.
Let \(\sigma \) be a state tape whose cell \(0\) reads the marker bit \(0\) and none of whose cells at positive natural positions reads \(0\) (they may be \(1\) or blank). Let \(q\) be an interpreter control state such that, for every combination of reads whose state-tape read is \(0\), the transition at \(q\) is the administrative action passing to a fixed state \(q'\) and moving the state cursor right, and for every combination whose state-tape read is not \(0\) it is the administrative action staying in \(q\) and moving the state cursor left, in both cases touching nothing else. Then for every natural number \(j\), starting from the evaluation configuration in state \(q\) with state tape \(\sigma \) and state cursor \(j\), exactly \(j + 1\) interpreter steps reach state \(q'\) with state cursor \(1\), leaving \(\sigma \), the table cursor, and all other data unchanged. The hypotheses are deliberately independent of whether the traversed cells are erased blanks or retained unary symbols.
Consider the table interpreter in its table-rewind phase (carrying a Boolean flag and a record index that it merely preserves), in an evaluation configuration over a table \(T\) with table cursor \(j - 1\) for some \(j \le |T|\), and with an arbitrary state tape \(\sigma \) and state cursor \(s\). Then exactly \(j + 1\) interpreter steps bring it to the count-parsing phase with the same flag and index and with the table cursor at \(0\); the state tape, state cursor, and all other data are unchanged.
Let \(q\) be an interpreter control state that, on every combination of reads with table read \(1\), stays in \(q\), writes a \(1\) at the state cursor, and moves both the table and state cursors right; and on every combination with table read \(0\) passes to a state \(q'\) with a fixed table movement \(d \in \{ -1, 0, +1\} \) and no other change. Suppose the table is \(l \cdot 1^n \cdot 0 \cdot r\), the state tape is the intact unary encoding of an index \(j\), the table cursor is \(|l|\), and the state cursor is \(j + 1\). Then exactly \(n + 1\) interpreter steps reach \(q'\) with table cursor \(|l| + n + d\), state tape the intact unary encoding of \(j + n\), and state cursor \(j + n + 1\); all other data are unchanged. This single gadget serves both initial-state extraction and live successor-state replacement.
The eight fixed bits of the serialized record of a one-work-tape machine action, as a function on indices \(0, \dots , 7\): bits \(0, 1\) encode the input-head movement, bits \(2, 3\) the optional work-tape write, bits \(4, 5\) the work-head movement, and bits \(6, 7\) the optional output emission, each in the two-bit encodings of the canonical record grammar.
For every one-work-tape machine action \(a\) over \(\{ 0,1\} \), the eight fixed bits of its serialized record decode back to its components: bits \(0, 1\) decode to \(a\)’s input-head movement, bits \(2, 3\) to its optional work-tape write, bits \(4, 5\) to its work-head movement, and bits \(6, 7\) to its optional output emission (bit \(6\) flags whether an output is emitted and bit \(7\) is the emitted bit).
For a coded machine \(M\) and a state \(q\) of \(M\), the list of the nine transition actions of \(q\): for each input read and each work read, ranging over blank, \(0\), \(1\) in the fixed serialization order with the input read varying slowest, the action \(M\) performs at \(q\) on that read pair.
For every coded machine \(M\), state \(q\) of \(M\), and pair of reads (an input symbol and a work symbol, each blank, \(0\), or \(1\)): the list of the nine actions of \(q\) has length \(9\), and its entry at the record index of the read pair (three times the serialization index of the input read plus the serialization index of the work read) is exactly the action \(M\) performs at \(q\) on those reads.
For a target control state \(q\), a table-cursor movement \(d_t \in \{ -1, 0, +1\} \), and a state-tape instruction consisting of an optional write and a cursor movement \(d_s\), the four-tape machine action that keeps the input head fixed, writes nothing and stays put on the simulated-work and marker tapes, moves the table cursor by \(d_t\), performs the optional write and the movement \(d_s\) on the state tape, emits no output, and passes to control state \(q\). By default \(d_t = 0\) and the state tape is untouched.
Applying the administrative action with target control state \(q'\), table movement \(d_t\), optional state-tape write \(w\), and state movement \(d_s\) to an interpreter evaluation configuration with table cursor \(p\) and state tape \(\sigma \) with cursor \(s\) yields the evaluation configuration over the same base with control state \(q'\), table cursor \(p + d_t\), state tape equal to \(\sigma \) updated at \(s\) by \(w\) when a write is present (and \(\sigma \) otherwise), and state cursor \(s + d_s\). The input head, the other two work tapes, and the accumulated output are unchanged.
The code-dependent block budget of the universal machine for an effective machine-coding scheme on a code string \(\alpha \):
where \(L\) is the length of the canonical serialization of the machine decoded from \(\alpha \) and \(N\) is that machine’s state count.
Let \(c\) be an effective machine-coding scheme with canonizer time bound \(T_c\), and let \(\alpha \) and \(x\) be bit strings. After exactly \(3\left\lvert \alpha \right\rvert +4+T_c(\left\lvert \alpha \right\rvert )\) transitions on the self-delimiting pairing \(\langle \alpha , x\rangle \), the buffered composition of the doubled-bit prefix parser with the canonizer of \(c\) has halted, its output is the serialization of the canonical transition table of the machine that \(c\) decodes from \(\alpha \), and its input head is at position \(2\left\lvert \alpha \right\rvert +3\). In particular this time depends only on the code, not on the suffix \(x\).
Let \(c\) be an effective machine-coding scheme, \(\alpha \) and \(x\) bit strings, and \(t\) a natural number. Then there is a Boolean boundary flag \(b\) such that, after \(3\left\lvert \alpha \right\rvert +4+t\) transitions on the self-delimiting pairing \(\langle \alpha , x\rangle \), the buffered composition of the doubled-bit prefix parser with the canonizer of \(c\) is in the second-phase buffered configuration embedding the configuration reached by the canonizer after \(t\) steps on input \(\alpha \), with flag \(b\) and with the physical input head still parked at position \(2\left\lvert \alpha \right\rvert +3\).
Let \(c\) be an effective machine-coding scheme and let \(\alpha \) and \(x\) be bit strings. Started on the self-delimiting pairing \(\langle \alpha , x\rangle \), the buffered composition of the doubled-bit prefix parser with the canonizer of \(c\) reaches, after exactly \(3\left\lvert \alpha \right\rvert +4\) transitions, the second-phase buffered configuration in which the canonizer stands at its initial configuration on input \(\alpha \), the virtual-boundary flag is set, and the physical input head is parked at position \(2\left\lvert \alpha \right\rvert +3\), on the first cell of the verbatim suffix. This step count does not depend on \(x\).
Given an effective machine-coding scheme — a representation of machines by bit strings equipped with a canonizer machine that computes, from any code, the serialization of the canonical transition table of the decoded machine within a stated time bound — this is the buffered composition of the doubled-bit prefix parser with that canonizer: the word emitted by the parser (the extracted code) is stored in a buffer, and the canonizer then runs with its input served from that buffer while the physical input head stays put.
For every machine \(M\), four-tape interpreter \(D\), and input \(x\), the initial configuration of the output-capture wrapper of \(M\) and \(D\) on \(x\) coincides with the first-phase embedding of the initial configuration of \(M\) on \(x\): blank table tape, blank interpreter tapes, and empty output.
Let \(M\) be a finite-tape machine over \(\{ 0,1\} \). For every four-tape configuration of the table interpreter, every assignment of contents and head positions to \(M\)’s work tapes, and every number of steps \(t\): running the combined capture machine (which pairs \(M\) with the fixed table interpreter) for \(t\) steps from the right-block embedding of the interpreter configuration equals the right-block embedding of the interpreter’s own \(t\)-step run, with the inactive tapes and heads unchanged.
Let \(c\) be a configuration of a machine \(M\) on input \(x\) and let \(t\) be a natural number such that \(M\) has not halted after any \(s \lt t\) steps from \(c\). Then \(t\) transitions of the output-capture wrapper of \(M\) and a four-tape interpreter \(D\), started from the first-phase embedding of \(c\), produce exactly the first-phase embedding of the configuration of \(M\) after \(t\) steps from \(c\).
Let \(M\) be a machine, \(D\) a four-tape interpreter, \(x\) an input, and \(T\) a natural number such that \(M\) has halted within \(T\) steps on \(x\). Then there exists \(t \le T+1\) such that, after exactly \(t\) transitions on \(x\), the output-capture wrapper of \(M\) and \(D\) stands in the interpreter-entry configuration corresponding to the configuration of \(M\) after \(T\) steps: the control is the interpreter’s initial state, the full halted configuration of \(M\) — including its parked input head — is preserved, and the output of \(M\) is captured on the table tape.
Let \(c\) be a configuration of a machine \(M\) on input \(x\) whose control state is live (not halted). Then one transition of the output-capture wrapper of \(M\) and a four-tape interpreter \(D\), applied to the first-phase embedding of \(c\), equals the first-phase embedding of the configuration obtained from \(c\) by one transition of \(M\). In particular any bit emitted by that transition of \(M\) — including a bit emitted on \(M\)’s halting transition — is appended on the table tape, and the wrapper itself remains live.
Let \(c\) be a halted configuration of a machine \(M\) on input \(x\). Then one transition of the output-capture wrapper of \(M\) and a four-tape interpreter \(D\), applied to the first-phase embedding of \(c\), yields the interpreter-entry configuration for \(c\): every component of \(c\) is preserved and the control becomes the interpreter’s initial state.
For a configuration \(c\) of a machine \(M\) with \(k\) work tapes on input \(x\), the corresponding first-phase configuration of the output-capture wrapper of \(M\) and an interpreter \(D\): the control records phase one together with \(c\)’s control state, the input head and the first \(k\) work tapes are those of \(c\), the table tape holds the word output by \(c\) so far (written from coordinate \(0\)) with its head one cell past the end of that word, the remaining three tapes are blank with heads at \(0\), and the wrapper’s real output is empty.
Let \(M\) be a finite-tape machine over \(\{ 0,1\} \), \(D\) a four-tape follower machine, and consider a configuration of \(M\) on an input \(x\). The endpoint configuration of the capture stage — which copies \(M\)’s accumulated output onto a fresh tape and hands control to \(D\) — equals the right-block embedding into the combined machine of the \(D\)-configuration with control at \(D\)’s initial state, the same physical input head as the source, first work tape holding the buffer tape of the captured output string with head at its right blank, and three blank tapes with heads at \(0\), where \(M\)’s work tapes and head positions are retained unchanged as the inactive left block.
For a configuration \(c\) of a machine \(M\) on input \(x\), the configuration of the output-capture wrapper of \(M\) and a four-tape interpreter \(D\) that agrees with the first-phase embedding of \(c\) in every component except that its control state is the interpreter’s initial state. The halted source configuration is thereby preserved as inactive data, with its output stored on the table tape.
Given a machine \(M\) over the binary alphabet with \(k\) work tapes and a four-tape machine \(D\) (the interpreter), the wrapper machine with \(k+4\) work tapes and two-phase control. In the first phase it runs \(M\) on the first \(k\) work tapes and the real input, except that every bit \(M\) emits is written onto a dedicated table tape — the first of the four extra tapes — instead of the real output. When \(M\) halts, a single administrative transition passes control to \(D\), which then runs on the four extra tapes, reads the original physical input, and emits output for real; the three extra tapes other than the table tape start blank.
The finite control-state type of the universal machine’s table interpreter: a fixed collection of phases (table rewind, count parsing, initial-state copying and skipping, state rewind, main dispatch, record skipping and reading, next-state copying, record application), each carrying only bounded registers — Boolean flags, record indices below \(9\), a field index below \(8\), and a buffer of eight action bits — together with a sink state for malformed reads. In particular the control is finite and independent of the number of states of the simulated machine; the unbounded simulated state is represented only on a tape.
Consider an interpreter evaluation configuration whose table tape is the buffer tape of a bit string \(T\) with cursor \(p\), whose state tape is \(\sigma \) with cursor \(s\), and whose remaining data come from a base configuration. The four symbols under its work heads are, in order: the symbol of the buffer tape of \(T\) at \(p\), the state symbol \(\sigma (s)\), and the two symbols read by the base configuration on its third and fourth work tapes.
Consider an evaluation configuration of the table interpreter with control state \(q\), table \(T\) with cursor \(p\), and state tape \(\sigma \) with cursor \(s\) over some base configuration. If the interpreter’s transition function at \(q\), applied to the physical input symbol and the represented reads (the buffer-tape symbol of \(T\) at \(p\), the state symbol \(\sigma (s)\), and the base configuration’s two remaining work reads), is the administrative action with target \(q'\), table movement \(d_t\), optional state write \(w\), and state movement \(d_s\), then one interpreter step takes the configuration to the evaluation configuration with control state \(q'\), table cursor \(p + d_t\), state tape updated by \(w\) at \(s\), and state cursor \(s + d_s\), leaving all other fields unchanged.
Given a base four-tape configuration (supplying the physical input position, the simulated work tape, the boundary-marker tape, and the accumulated output), a control state \(q\), a bit string \(T\) with a table cursor \(p \in \mathbb {Z}\), and a state tape \(\sigma \colon \mathbb {Z} \to \{ 0, 1, \mathrm{blank}\} \) with a state cursor \(s\): the configuration in control state \(q\) whose first work tape is the buffer tape of \(T\) with head at \(p\), whose second work tape is \(\sigma \) with head at \(s\), and which agrees with the base configuration in every other field.
Given four values \(a, b, c, d\) of a common type, the function on the four-element index type sending index \(0\) to \(a\), index \(1\) to \(b\), index \(2\) to \(c\), and index \(3\) to \(d\). It packages one entry per work tape of a four-tape machine into a single function, without a variable-size register.
The header of the canonical serialization of a coded machine \(M\) with state count \(m\) and initial state index \(q_0\): the pair encoding of the binary digit string of \(m\) with the empty string (each digit doubled, then the unequal delimiter pair), followed by \(1^{q_0} \cdot 0\), the initial state in unary with its terminator. It comprises everything preceding the transition records.
Let \(\alpha \) be a bit string, let \(c\) be a configuration of a machine with \(k\) work tapes on input \(x\) whose input head is at position \(p\), and let \(d \in \{ -1, 0, +1\} \) be a direction. Let \(b\) be the marker-derived boundary flag at \(p\) (false exactly at \(p = 0\)) and let \(m\) be the clamped virtual move determined by \(b\), the symbol under \(c\)’s input head, and \(d\), which suppresses stepping outward from a boundary blank. Then moving the physical head from the embedded position \(2\left\lvert \alpha \right\rvert +2+p\) on the self-delimiting pairing \(\langle \alpha , x\rangle \) by \(m\) lands exactly on the embedding of the position reached from \(p\) by the clamped move in direction \(d\), and, as integers, \(p + m\) equals that new virtual position.
Let \(\alpha \) be a bit string and let \(c\) be a configuration of a machine with \(k\) work tapes on input \(x\); embed \(c\) into the self-delimiting pairing \(\langle \alpha , x\rangle \) by shifting its input head by the suffix offset \(2\left\lvert \alpha \right\rvert +2\). Mask the embedded read as follows: if the marker tape carrying a single mark at coordinate \(0\) shows its mark at the coordinate equal to \(c\)’s virtual head position — that is, exactly at virtual position \(0\) — declare the read blank; otherwise take the symbol physically scanned by the embedded configuration. This masked read equals the input symbol that \(c\) scans natively on \(x\), including at the left and right boundary blanks (and for an empty suffix).
For every head position \(p \in \{ 0, 1, \dots , n+1\} \) on an input of length \(n\) (with its two boundary blanks), the Boolean flag “the marker tape carrying a single mark at coordinate \(0\) does not show its mark at coordinate \(p\)” satisfies the boundary-tag contract of the clamped virtual-move operation: the flag is false at the left boundary \(p = 0\) and true at the right blank \(p = n+1\), including when the input is empty.
Given a bit string \(\alpha \) and a configuration of a machine with \(k\) work tapes on input \(x\), this forms the configuration on the self-delimiting pairing \(\langle \alpha , x\rangle \) that keeps the control state, work tapes, work-head positions, and output unchanged and translates only the input-head position by the offset \(2\left\lvert \alpha \right\rvert +2\) of the verbatim suffix.
Given bit strings \(\alpha \) and \(x\), this maps a virtual input-head position \(p \in \{ 0, 1, \dots , \left\lvert x\right\rvert +1\} \) on the suffix \(x\) (including both boundary blanks) to the physical head position \(2\left\lvert \alpha \right\rvert +2+p\) on the self-delimiting pairing \(\langle \alpha , x\rangle \). Virtual position \(0\) lands on the separator’s last cell.
The fixed four-tape machine over \(\{ 0,1\} \) that interprets a serialized transition table. Its four work tapes hold the captured program table, the current simulated state in unary behind a permanent origin marker, the simulated machine’s work tape, and a marker tape recording whether the virtual input head sits at the simulated left boundary. To perform one simulated transition it rewinds the table, skips nine records per unary state symbol while consuming the unary state destructively, selects within the reached block the record indexed by the current simulated reads, buffers its eight fixed action bits, copies the successor-state unary field back onto the state tape, and applies the decoded action, adjusting the input-head movement at the virtual left boundary. The control is finite and independent of the number of coded states; malformed administrative reads enter a non-halting sink state.
One step of the table interpreter from its entry configuration on a captured table \(T\) installs the two permanent markers — the origin marker \(0\) on the state tape and the boundary bit \(1\) on the marker tape — and starts the unconditional table rewind: the result is the evaluation configuration over the initialization base whose control is the initial table-rewind phase, with table cursor \(|T| - 1\), the unary state tape encoding the index \(0\), and state cursor \(1\).
Suppose the captured table \(T\) is the pair encoding of a bit string \(\beta \) (the count field) with a string of the form \(1^n \cdot 0 \cdot \rho \) (the initial state \(n\) in unary, followed by an arbitrary remainder \(\rho \)). Then from the interpreter’s entry configuration with physical input head parked at a position \(p\), exactly \(|T| + 2|\beta | + 2n + 7\) interpreter steps reach the evaluation configuration over the initialization base at \(p\) whose control is the main dispatch state, with table cursor \(2|\beta | + 2 + n + 1\) (just past the serialized header), the unary state tape encoding the index \(n\), and state cursor \(1\). No transition-table lookup is involved yet.
The inactive data carried during interpreter initialization, packaged as a base configuration: control at the interpreter’s main dispatch state, physical input head parked at a given position \(p\), table and state tapes blank with heads at \(0\), simulated work tape blank with head at \(0\), the boundary-marker tape holding a single \(1\) in cell \(0\) with its head at position \(1\) (also correct for empty suffixes), and no output emitted.
The interpreter’s entry configuration for a captured table \(T\): control at the interpreter’s start state, physical input head at a given position \(p\) (already parked at the start of the input suffix), first work tape holding the buffer tape of \(T\) with its head at \(|T|\) (the right blank of the table), the other three work tapes blank with heads at \(0\), and no output emitted.
The number of interpreter steps needed to decode the successor field of a serialized record and install the successor state in unary, before the decoded action is applied: \(1\) when there is no successor state (the action halts), and \(2q + 4\) when the successor state has index \(q\).
The number of \(1\)s in the self-delimiting unary successor field of a serialized record: \(0\) when there is no successor state (the action halts), and \(q + 1\) when the successor state has index \(q\) (one \(1\) for the live flag and \(q\) for the index).
Let \(\alpha \) and \(x\) be bit strings and let \(j \le \left\lvert \alpha \right\rvert \). Started on the self-delimiting pairing \(\langle \alpha , x\rangle \) (each bit of \(\alpha \) doubled, then the separator pair \(01\), then \(x\) verbatim), the doubled-bit prefix parser is, after exactly \(2j\) transitions, in the “expecting a first half” state with its input head at position \(2j+1\) — head positions count the left boundary blank as position \(0\), so the head stands on the cell of index \(2j\) — having emitted precisely the first \(j\) bits of \(\alpha \).
For all bit strings \(\alpha \) and \(x\), after exactly \(2\left\lvert \alpha \right\rvert +1\) transitions on the self-delimiting pairing \(\langle \alpha , x\rangle \), the doubled-bit prefix parser is in the live state recording the separator’s first bit \(0\), its input head is at position \(2\left\lvert \alpha \right\rvert +2\), and it has emitted exactly \(\alpha \).
For all bit strings \(\alpha \) and \(x\), the doubled-bit prefix parser, started on the self-delimiting pairing \(\langle \alpha , x\rangle \), halts after exactly \(2\left\lvert \alpha \right\rvert +2\) transitions with output exactly \(\alpha \) and with its input head parked at position \(2\left\lvert \alpha \right\rvert +3\), that is, on the first cell of the verbatim suffix \(x\) (the right blank when \(x\) is empty). Both the step count and the emitted word are independent of \(x\).
Let the doubled-bit prefix parser stand, on the self-delimiting pairing \(\langle \alpha , x\rangle \) of bit strings \(\alpha \) and \(x\), in a live control state \(q\) at input-head position \(p\) with output word \(w\) emitted so far, and suppose the symbol scanned at \(p\) is \(b\). Then one transition produces again a configuration with no work-tape content: its control state and head movement are those prescribed by the parser’s transition function on \((q, b)\), and the bit emitted by that transition (if any) is appended to \(w\).
Given bit strings \(\alpha \) and \(x\), a control state, an input-head position on the self-delimiting pairing \(\langle \alpha , x\rangle \), and an output word, this assembles the corresponding full configuration of the doubled-bit prefix parser; since that parser has no work tapes, the work-tape and work-head components are vacuous.
A Turing machine over the binary alphabet with no work tapes and two kinds of live control states: “expecting the first half of a doubled pair” and “having just read a first half \(b\)”. In the expecting state it records the scanned bit and moves right (halting if it scans a blank); in the state remembering \(b\) it compares the scanned cell with \(b\) — on a match it emits the single bit \(b\), moves right, and returns to the expecting state, while on a mismatch (which on a valid paired encoding occurs exactly at the separator’s second bit) it moves right and halts without emitting.
The bit-string serialization of one transition record of a one-work-tape machine action over \(\{ 0,1\} \): two bits encoding the input-head movement, two bits encoding the optional work-tape write, two bits encoding the work-head movement, two bits encoding the optional output emission, and finally the optional successor state as a self-delimiting unary field — a single \(0\) when the action halts, and \(1 \cdot 1^{q} \cdot 0\) when the successor state has index \(q\).
For a pair of read symbols — an input symbol and a work symbol, each blank, \(0\), or \(1\) — the index in \(\{ 0, \dots , 8\} \) of the corresponding transition record within the nine-record block of one coded state: three times the serialization index of the input read (blank \(\mapsto 0\), bit \(0 \mapsto 1\), bit \(1 \mapsto 2\)) plus the serialization index of the work read.
The canonical transition table of a coded machine \(M\): the concatenation, over every state of \(M\) in order and over every pair of reads (input read, work read) in the fixed serialization order blank, \(0\), \(1\), of the serialized transition record of the action \(M\) performs at that state and read pair — nine records per state, including all nine read pairs at every live state.
For an effective machine-coding scheme \(c\), a code string \(\alpha \), and an input \(x\): the relation between a configuration \(s\) of the machine decoded from \(\alpha \) (in one-work-tape form) on \(x\) and a configuration \(t\) of the complete universal-machine candidate on the pair encoding \(\langle \alpha , x \rangle \) that holds when there exist a table cursor \(p\) not exceeding the length of the canonical serialization of the decoded machine, together with contents and head positions for the inactive canonizer tapes, such that \(t\) is the right-block embedding, over those inactive tapes, of the interpreter checkpoint of \(s\) with table cursor \(p\).
Let \(c\) be an effective machine-coding scheme, \(\alpha \) a code string, and \(x\) an input. If the concrete checkpoint relation holds between a configuration \(s\) of the machine decoded from \(\alpha \) on \(x\) and a configuration \(t\) of the complete universal-machine candidate on the pair encoding \(\langle \alpha , x \rangle \), then \(s\) has halted if and only if \(t\) has halted.
Let \(c\) be an effective machine-coding scheme, \(\alpha \) a code string, and \(x\) an input. If the concrete checkpoint relation holds between a configuration \(s\) of the machine decoded from \(\alpha \) on \(x\) and a configuration \(t\) of the complete universal-machine candidate on the pair encoding \(\langle \alpha , x \rangle \), then the accumulated output of \(s\) equals the accumulated output of \(t\).
For every effective machine-coding scheme \(c\), code string \(\alpha \), and input \(x\), there is a time \(t\) not exceeding the startup bound for \(\alpha \) such that the concrete checkpoint relation holds between the initial configuration of the machine decoded from \(\alpha \) on \(x\) and the configuration of the complete universal-machine candidate after \(t\) steps from its initial configuration on the pair encoding \(\langle \alpha , x \rangle \).
For a coded machine \(M\), a code string \(\alpha \), a configuration of \(M\)’s one-work-tape form on an input \(x\), and a table-cursor value: the four-tape interpreter configuration on the pair-encoded input \(\langle \alpha , x \rangle \) whose control is the main dispatch state when the source is live and halted when the source has halted; whose physical input head represents the source input head, offset past the encoded prefix; whose table tape holds the buffer tape of the canonical serialization of \(M\) at the given cursor (the only administrative coordinate not determined by the source); whose state tape holds the source’s current state index in unary with state cursor \(1\); whose third work tape mirrors the source work tape with the same head; whose marker tape holds a single \(1\) at the origin with its head at the source input position (so it reads \(1\) exactly at the simulated left boundary); and whose accumulated output equals the source output.
The interpreter control state entered after the last record of a skip request: the record-group phase with index \(i\) when the destination carries an index \(i\), and the action reader (starting at field \(0\) with a cleared bit buffer) when the destination is empty.
The startup time bound of the universal machine for a coding scheme on a code string \(\alpha \):
where \(t_c\) is the scheme’s canonizer time bound, \(s\) is the canonical serialization of the machine decoded from \(\alpha \), \(\ell \) is the number of binary digits of that machine’s state count, and \(q_0\) is its initial state index. It depends on \(\alpha \) but not on the input suffix, and includes complete interpreter setup.
The state-tape contents during a destructive table lookup, parametrized by the number \(c\) of unary cells already erased and the number \(m\) still remaining: cell \(0\) holds the permanent origin marker \(0\), cells \(c + 1, \dots , c + m\) hold \(1\)s, and all other cells are blank.
The universal-machine candidate associated with an effective machine-coding scheme: the composite finite-tape machine over \(\{ 0,1\} \) that first virtually runs the scheme’s canonizer to extract the code prefix of the pair-encoded input and produce the canonical serialized table, then captures that table onto a work tape, and finally hands control to the fixed four-tape table interpreter.
For every action of a \(k\)-tape machine, every configuration, and every work tape index \(i \lt k\), the head position of tape \(i\) after applying the action differs from its position before by at most one: \(\left\lvert p' - p\right\rvert \le 1\), where \(p\) and \(p'\) are the positions of the head of tape \(i\) before and after the action is applied.
Work in the symplectic space \(V = \mathbb {F}_p^{\, n} \times \mathbb {F}_p^{\, n}\) over the prime field \(\mathbb {F}_p\), equipped with the symplectic form \(\omega \), and let \(S \le V\) be an isotropic submodule, so that \(\omega (u,v) = 0\) for all \(u, v \in S\). Suppose \(A, B \subseteq \{ 0, 1, \dots , n-1\} \) are two disjoint sets of coordinates, each correctable for \(S\): every vector of \(S^{\perp _\omega }\) supported on \(A\) lies in \(S\), and likewise every vector of \(S^{\perp _\omega }\) supported on \(B\) lies in \(S\). Then the logical dimension \(k(S) = n - \dim S\) satisfies
Fix natural numbers \(k, n, d\) with \(k \ge 1\), \(n \ge 2\), and \(d \ge 1\). Then there is a positive constant \(C_0 \gt 0\) such that for every natural number \(n_P\) satisfying \(n_P\, d \le n k\), one has
where \(\log \) denotes the natural logarithm.
For parameters \(k, n, d\) and a message \(b : \mathrm{Fin}\, k \to \mathbb {Z}\), the modelled value of the \(2\)-XOR polynomial \(g_b\) obtained from the decomposition, namely the real number
The definition is abstract: the quantity depends only on \(n\), \(k\) and the threshold \(d\), and is constant in the edge count and in \(b\).
Fix a prime \(p\) and an integer \(n\), and let \(V = \mathbb {F}_p^n \times \mathbb {F}_p^n\). Taking \(C\) to be the full index set \(\{ 0,1,\dots ,n-1\} \), the support submodule \(V_C = \{ v \in V \mid \mathrm{supp}(v) \subseteq C\} \) is all of \(V\); that is, \(V_C = V\).
Let \(\omega \) be a \(k \times d\) random array whose entries are independent and identically distributed with the Rademacher distribution \(\tfrac {1}{2}\delta _{-1} + \tfrac {1}{2}\delta _{+1}\), taking each of the values \(-1\) and \(+1\) with probability \(1/2\). Then for every index \((i,j)\) with \(i \in \{ 0,\dots ,k-1\} \) and \(j \in \{ 0,\dots ,d-1\} \), the single entry \(\omega _{ij}\) has variance
Let \(X\) be a real random variable distributed according to the Rademacher distribution \(\mathrm{Rad} = \tfrac 12\delta _{-1} + \tfrac 12\delta _{+1}\), so that \(X\) takes the values \(-1\) and \(+1\) each with probability \(1/2\). Then the variance of \(X\) is
Fix an integer \(k \ge 1\), a vector \(x \in \mathbb {R}^d\), and a row index \(i \in \{ 1,\dots ,k\} \). Let \(A\) be the Rademacher matrix drawn from the joint Rademacher product measure, so its \(k \cdot d\) entries \(A_{ij} = \omega _{ij}/\sqrt{k}\) are independent, each equal to \(\pm 1/\sqrt{k}\) with probability \(1/2\). Then the \(i\)-th coordinate of the image \(Ax\) has variance
Let \(k \gt 0\) and let \(x \in \mathbb {R}^d\). Draw a \(k \times d\) array \(\omega \) whose entries \(\omega _{ij}\) are independent Rademacher variables, each equal to \(-1\) or \(+1\) with probability \(\tfrac {1}{2}\). Then for any indices \(i \in \{ 1,\dots ,k\} \) and \(j \in \{ 1,\dots ,d\} \), the real random variable \(\omega \mapsto (x_j/\sqrt{k})\, \omega _{ij}\) has variance
Let \(G\) be a finite two-player zero-sum game with \(M \ge 1\) row actions and \(N \ge 1\) column actions, given by a payoff matrix \(A\) whose entries satisfy \(A_{ij} \in [0,1]\), where the row player maximises and the column player minimises. Let \(p = (p_i)\) be a mixed strategy for the row player and \(q = (q_j)\) a mixed strategy for the column player. Then the row player’s guaranteed payoff under \(p\) never exceeds the payoff the column player can be forced above under \(q\):
Let \(p_1,\dots ,p_n\) be a probability vector, that is \(p_i \ge 0\) for every \(i\) and \(\sum _{i} p_i = 1\), and let \(\ell _1,\dots ,\ell _n \in \mathbb {R}\) satisfy \(0 \le \ell _i \le 1\) for every \(i\). Then for every \(\eta \in \mathbb {R}\),
Let \(\beta \in (0,1)\) and let \(n \ge 1\) be an integer, and let \(M^\ast , M_{\mathrm{WM}}\) be nonnegative integers satisfying
Then
where \(\log \) denotes the natural logarithm.
Let \(n \ge 2\) be an integer and let \(\beta \in (0,1)\). Suppose \(M^*\) and \(M_{\mathrm{WM}}\) are nonnegative integers satisfying
Then
Let \(\beta \in (0,1)\), let \(W : \mathbb {N}\to \mathbb {R}\) be a sequence of positive reals, and let \(M\) be a natural number such that \(W(k+1) \le W(k) \cdot \tfrac {1+\beta }{2}\) for every \(k \lt M\). Then
Fix an integer \(n \ge 1\) and a real number \(\beta \) with \(0 \lt \beta \lt 1\), and let \(W \colon \mathbb {N}\to \mathbb {R}\) be a sequence of strictly positive reals with \(W(0) = n\). Suppose that for every \(k \lt M\) the sequence shrinks by at least the factor \((1+\beta )/2\) at each step, in the sense that
Then
A finite two-player zero-sum game with \(M\) row actions and \(N\) column actions is a record consisting of a payoff matrix \(A : \mathrm{Fin}\, M \to \mathrm{Fin}\, N \to \mathbb {R}\) together with proofs that every entry satisfies \(0 \le A_{ij} \le 1\). The row player seeks to maximise the payoff; the column player seeks to minimise it.
Given a game \(G\) and a sequence of column responses \(j_0, \dots , j_{T-1}\), the induced loss sequence for the row player is \(\ell _t(i) = 1 - A_{i,j_t} \in [0,1]\). This translates a zero-sum game into the abstract loss-sequence framework used by Hedge.
Let \(G\) be a finite two-player zero-sum game with \(M\) row actions and \(N\) column actions, given by a payoff matrix \(A\) whose entries satisfy \(0 \le A_{ij} \le 1\), and let \(j_0, \dots , j_{T-1}\) be any sequence of column responses. Then the induced loss sequence for the row player, defined by \(\ell _t(i) = 1 - A_{i,j_t}\) for each round \(t \lt T\) and row action \(i\), is valid: every loss lies in \([0,1]\), i.e. \(0 \le \ell _t(i) \le 1\) for all \(t\) and \(i\).
Let \(k \ge 1\), and let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\). For each \(j \in \mathbb {Z}_k\) let \(\widetilde{f}_j(x) = \omega _k^{\, j\, f(x)}\) denote the \(j\)-twisted lift of \(f\), and write \(\widehat{\widetilde{f}_j}(s)\) for its Fourier coefficient at frequency \(s \in \mathbb {Z}_k^n\). Then the BLR acceptance probability of \(f\) admits the representation
where \(|\cdot |\) is the complex modulus and \(\operatorname {Re}\) the real part.
Let \(k \ge 1\) and \(n \ge 0\), and let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) be a linear function, meaning \(f(x+y) = f(x) + f(y)\) for all \(x, y \in \mathbb {Z}_k^n\). Then the BLR acceptance probability of \(f\) equals \(1\); that is,
where \(x\) and \(y\) range independently and uniformly over \(\mathbb {Z}_k^n\).
Let \(p\) be prime and fix \(n \ge 0\). Let \(f : \mathbb {Z}_p^n \to \mathbb {Z}_p\) be a normalized function (so \(f(0)=0\)), and let \(0 \le \varepsilon \le 1\) be such that \(f\) is \(\varepsilon \)-far from every linear map, meaning \(\mathrm{dist}(f,g) \ge \varepsilon \) for every linear \(g : \mathbb {Z}_p^n \to \mathbb {Z}_p\), where \(\mathrm{dist}(f,g) = \Pr _x[f(x) \ne g(x)]\). Then the BLR acceptance probability of \(f\) satisfies
where \(x,y\) range uniformly and independently over \(\mathbb {Z}_p^n\).
Let \(k \ge 2\), and let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) be normalized, so that \(f(0) = 0\). Fix a real number \(\varepsilon \) with \(0 \le \varepsilon \le 1\), and suppose \(f\) is \(\varepsilon \)-far from linear: for every linear function \(g : \mathbb {Z}_k^n \to \mathbb {Z}_k\) (one satisfying \(g(x+y) = g(x)+g(y)\) for all \(x,y\)), the fraction of inputs on which \(f\) and \(g\) disagree is at least \(\varepsilon \). Then the acceptance probability of the BLR linearity test, taken over independent uniform \(x, y \in \mathbb {Z}_k^n\), obeys
where \(\varphi \) denotes Euler’s totient function.
Let \(p\) be prime and let \(f : \mathbb {Z}_p^n \to \mathbb {Z}_p\) be normalized, meaning \(f(0) = 0\). Fix \(\varepsilon \) with \(0 \le \varepsilon \le 1\), and suppose \(f\) is \(\varepsilon \)-far from linear: every linear function \(g : \mathbb {Z}_p^n \to \mathbb {Z}_p\) (one satisfying \(g(x+y) = g(x)+g(y)\) for all \(x,y\)) has \(\operatorname{dist}(f,g) \ge \varepsilon \), where \(\operatorname{dist}(f,g) = \Pr _x[f(x) \ne g(x)]\). Then the BLR acceptance probability of \(f\) obeys
the probability being taken over \(x,y\) drawn independently and uniformly from \(\mathbb {Z}_p^n\).
Let \(p\) be a prime, let \(n \ge 0\), and let \(f : \mathbb {Z}_p^n \to \mathbb {Z}_p\) be a function, whose lift to roots of unity is \(\widetilde{f}(x) = \omega _p^{f(x)}\) with \(\omega _p = e^{2\pi i/p}\). Write \(\widehat{\widetilde{f}}(s)\) for the Fourier coefficient of \(\widetilde{f}\) at frequency \(s \in \mathbb {Z}_p^n\). If \(A \in \mathbb {R}\) is such that \(\bigl\| \widehat{\widetilde{f}}(s)\bigr\| \le A\) for every \(s \in \mathbb {Z}_p^n\), then
A function \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) satisfies ZkBLR.epsilon_far_from_linear_normalized with parameter \(\varepsilon \) if \(f\) is normalized (\(f(0)=0\)), \(0\le \varepsilon \le 1\), and \(\mathrm{dist}(f,g)\ge \varepsilon \) for every linear \(g\).
Let \(k \ge 1\) and write \(\omega _k = e^{2\pi i/k}\) for the primitive \(k\)-th root of unity. For each \(a \in \mathbb {Z}/k\mathbb {Z}\),
where each exponent \(ja \in \mathbb {Z}/k\mathbb {Z}\) is taken via its canonical representative in \(\{ 0, 1, \ldots , k-1\} \).
Let \(k \ge 1\) and write \(\omega _k = e^{2\pi i/k}\) for the primitive \(k\)-th root of unity. For every \(a \in \mathbb {Z}/k\mathbb {Z}\),
where the exponent \(ja\) is computed in \(\mathbb {Z}/k\mathbb {Z}\) and then lifted to its canonical representative in \(\{ 0,\ldots ,k-1\} \), and the left-hand side equals \(1\) when \(a = 0\) and \(0\) otherwise.
For \(j \in \mathbb {Z}_k\), the \(j\)-twisted lift of \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) is \(\widetilde{f}_j(x) = \omega _k^{j \cdot f(x)}\). The case \(j=1\) recovers ZkBLR.lift_omega; the case \(j=0\) gives the constant function \(1\).
Fix an integer \(k \ge 1\), and let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) be arbitrary. Write \(\widetilde{f}_0(x) = \omega _k^{\, 0 \cdot f(x)}\) for the \(0\)-twisted lift of \(f\) (equivalently, the constant function \(1\)), and let \(\widehat{\widetilde{f}_0}(s)\) denote its Fourier coefficient at frequency \(s \in \mathbb {Z}_k^n\). Then
where \(|\cdot |\) is the complex modulus.
Fix an integer \(k \ge 1\) and a dimension \(n\), and let \(f : (\mathbb {Z}/k\mathbb {Z})^n \to \mathbb {Z}/k\mathbb {Z}\) be any function. Then \(f\) is linear — that is, \(f(x+y) = f(x) + f(y)\) for all \(x, y \in (\mathbb {Z}/k\mathbb {Z})^n\) — if and only if there is a vector \(s \in (\mathbb {Z}/k\mathbb {Z})^n\) with \(f(x) = s \cdot x = \sum _{i=0}^{n-1} s_i x_i\) for every \(x\), all sums taken modulo \(k\).
Let \(k \ge 1\) and \(n\) be natural numbers, and let \(f : (\mathbb {Z}/k\mathbb {Z})^n \to \mathbb {Z}/k\mathbb {Z}\) be a linear function, meaning \(f(x+y) = f(x) + f(y)\) for all \(x, y \in (\mathbb {Z}/k\mathbb {Z})^n\). Then \(f\) is normalized: \(f(0) = 0\).
Let \(f : (\mathbb {Z}/k\mathbb {Z})^n \to \mathbb {Z}/k\mathbb {Z}\) be any function on the \(n\)-dimensional space over \(\mathbb {Z}/k\mathbb {Z}\), and let its normalization be \((\mathrm{normalize}\, f)(x) = f(x) - f(0)\). Then this normalization vanishes at the origin: \((\mathrm{normalize}\, f)(0) = 0\).
Let \(p\) be a prime and \(n \ge 0\), and write \(\omega _p = e^{2\pi i/p}\). For a function \(f : \mathbb {Z}_p^n \to \mathbb {Z}/p\mathbb {Z}\), let \(\widetilde{f} : \mathbb {Z}_p^n \to \mathbb {C}\) be its lift \(\widetilde{f}(x) = \omega _p^{f(x)}\), and let \(\widehat{\widetilde{f}}(s)\) denote its Fourier coefficient at frequency \(s \in \mathbb {Z}_p^n\). Then
Let \(k \ge 1\) and \(n \ge 0\) be integers, and write \(\omega _k = e^{2\pi i/k}\). Fix \(j \in \mathbb {Z}/k\mathbb {Z}\) and an arbitrary function \(f : \mathbb {Z}_k^n \to \mathbb {Z}/k\mathbb {Z}\), and let \(\widetilde{f}_j : \mathbb {Z}_k^n \to \mathbb {C}\) be the \(j\)-twisted lift given by \(\widetilde{f}_j(x) = \omega _k^{\, j\cdot f(x)}\), with Fourier coefficients \(\widehat{\widetilde{f}_j}(s)\) indexed by the frequencies \(s \in \mathbb {Z}_k^n\). Then
Fix an integer \(k \ge 2\) and let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) be \(\varepsilon \)-far from linear in the normalized sense: that is, \(f(0) = 0\), the parameter satisfies \(0 \le \varepsilon \le 1\), and \(\operatorname{dist}(f, g) \ge \varepsilon \) for every linear \(g : \mathbb {Z}_k^n \to \mathbb {Z}_k\). Write \(\widetilde{f}(x) = \omega _k^{f(x)}\) for the lift of \(f\) to the unit circle, where \(\omega _k = e^{2\pi i/k}\), and let \(\widehat{\widetilde{f}}(s)\) denote its Fourier coefficient at frequency \(s\). Then for every \(s \in \mathbb {Z}_k^n\),
Let \(p\) be a prime, let \(n \ge 0\), and write \(\omega _p = e^{2\pi i/p}\) for the primitive \(p\)-th root of unity. Let \(f : \mathbb {Z}_p^{\, n} \to \mathbb {Z}_p\) be normalized (\(f(0)=0\)) and \(\varepsilon \)-far from linear, meaning \(0 \le \varepsilon \le 1\) and \(\Pr _x[f(x)\ne g(x)] \ge \varepsilon \) for every additive homomorphism \(g : \mathbb {Z}_p^{\, n} \to \mathbb {Z}_p\). Fix \(j \in \mathbb {Z}_p\) with \(j \ne 0\), and for \(s \in \mathbb {Z}_p^{\, n}\) let \(\widehat{h}(s) = \mathbb {E}_x\bigl[h(x)\, \omega _p^{-\, s\cdot x}\bigr]\) denote the Fourier coefficient of a function \(h : \mathbb {Z}_p^{\, n} \to \mathbb {C}\) at frequency \(s\), where \(s\cdot x = \sum _{i} s_i x_i \in \mathbb {Z}_p\). Then the \(j\)-twisted lift \(\widetilde{f}_j(x) = \omega _p^{\, j\, f(x)}\) satisfies
for every \(s \in \mathbb {Z}_p^{\, n}\).
Let \(k \ge 2\) and let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) be normalized and \(\varepsilon \)-far from linear, meaning that \(f(0) = 0\), that \(0 \le \varepsilon \le 1\), and that \(\operatorname{dist}(f,g) \ge \varepsilon \) for every linear \(g : \mathbb {Z}_k^n \to \mathbb {Z}_k\) (where \(\operatorname{dist}\) is the fraction of inputs on which two functions disagree). Fix a unit \(j \in (\mathbb {Z}/k\mathbb {Z})^\times \) and a frequency \(s \in \mathbb {Z}_k^n\), and let \(\widetilde{f}_j(x) = \omega _k^{\, j\, f(x)}\) be the corresponding \(j\)-twisted lift, with \(\omega _k = e^{2\pi i/k}\). Then the Fourier coefficient of \(\widetilde{f}_j\) at \(s\) has real part bounded by
Let \(k \ge 2\) and let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\), with lift \(\widetilde{f}(x) = \omega _k^{f(x)}\) where \(\omega _k = e^{2\pi i/k}\). Then for every frequency \(s \in \mathbb {Z}_k^n\), the real part of the Fourier coefficient of \(\widetilde{f}\) at \(s\) is bounded above by
where \(\chi _s(x) = s\cdot x\) is the linear character and \(\mathrm{dist}(f,\chi _s)\) is the fraction of inputs \(x\) on which \(f(x) \ne s\cdot x\).
Let \(k \ge 2\) and write \(\omega _k = e^{2\pi i / k}\) for the primitive \(k\)-th root of unity. For every nonzero \(a \in \mathbb {Z}/k\mathbb {Z}\),
where \(\omega _k^{\, a}\) denotes \(\omega _k\) raised to the canonical representative of \(a\) in \(\{ 0, 1, \ldots , k-1\} \).
Fix \(k \ge 1\) and \(n \ge 0\), and work over the group \(\mathbb {Z}_k^n = (\mathbb {Z}/k\mathbb {Z})^n\) with primitive \(k\)-th root of unity \(\omega _k = e^{2\pi i/k}\). Given \(j \in \mathbb {Z}/k\mathbb {Z}\) and a map \(f \colon \mathbb {Z}_k^n \to \mathbb {Z}/k\mathbb {Z}\), let \(F(x) = \omega _k^{\, j\cdot f(x)}\) be its \(j\)-twisted lift, and for \(s \in \mathbb {Z}_k^n\) let \(\widehat{F}(s) = \mathbb {E}_x\! \big[F(x)\, \overline{\omega _k^{\, s\cdot x}}\big]\) be its Fourier coefficient, where \(s\cdot x = \sum _{i} s_i x_i \in \mathbb {Z}/k\mathbb {Z}\) and \(\mathbb {E}\) denotes the uniform average over \(\mathbb {Z}_k^n\). Then the triple-product expectation over independent uniform \(x,y\) equals a cubic sum of the Fourier coefficients:
Let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) be a normalized function that is \(\varepsilon \)-far from linear, meaning \(f(0) = 0\), \(0 \le \varepsilon \le 1\), and \(\operatorname{dist}(f, g) \ge \varepsilon \) for every linear \(g : \mathbb {Z}_k^n \to \mathbb {Z}_k\). Then for any unit \(j \in (\mathbb {Z}_k)^\times \), the scaled function \(x \mapsto j \cdot f(x)\) is again normalized and \(\varepsilon \)-far from linear.
Let \(p\) be prime, and let \(f : \mathbb {Z}_p^n \to \mathbb {Z}_p\) be normalized and \(\varepsilon \)-far from linear: that is, \(f(0) = 0\), the parameter satisfies \(0 \le \varepsilon \le 1\), and \(\mathrm{dist}(f,g) \ge \varepsilon \) for every linear \(g : \mathbb {Z}_p^n \to \mathbb {Z}_p\). Write \(\widetilde{f}(x) = \omega _p^{f(x)}\) for the lift of \(f\) to the \(p\)-th roots of unity, and \(\hat{\widetilde{f}}(s)\) for its Fourier coefficient at frequency \(s \in \mathbb {Z}_p^n\). Then
Let \(k \ge 1\) and \(n\) be integers, fix \(j \in \mathbb {Z}/k\mathbb {Z}\) and a function \(f : \mathbb {Z}_k^n \to \mathbb {Z}/k\mathbb {Z}\), and let \(\widetilde{f}_j(x) = \omega _k^{\, j\, f(x)}\) be its \(j\)-twisted lift, with Fourier coefficients \(\widehat{\widetilde{f}_j}(s)\) indexed by frequencies \(s \in \mathbb {Z}_k^n\). Then
where \(\lvert \cdot \rvert \) denotes the complex modulus and \(\operatorname {Re}\) the real part.
Let \(p\) be prime and let \(f : \mathbb {Z}_p^n \to \mathbb {Z}_p\) be normalized, meaning \(f(0) = 0\), and \(\varepsilon \)-far from linear: \(0 \le \varepsilon \le 1\) and \(\mathrm{dist}(f,g) \ge \varepsilon \) for every linear function \(g : \mathbb {Z}_p^n \to \mathbb {Z}_p\), where \(\mathrm{dist}\) denotes the fraction of inputs on which the two functions disagree. Fix a nonzero \(j \in \mathbb {Z}_p\) and let \(\widetilde{f}_j(x) = \omega _p^{\, j\, f(x)}\) be the associated \(j\)-twisted lift, with \(\omega _p = e^{2\pi i / p}\), and write \(\widehat{\widetilde{f}_j}(s)\) for its Fourier coefficient at frequency \(s\). Then
Let \(k \ge 2\), and let \(f : \mathbb {Z}_k^n \to \mathbb {Z}_k\) be normalized (that is, \(f(0) = 0\)) and \(\varepsilon \)-far from linear for some \(0 \le \varepsilon \le 1\), meaning that \(\mathrm{dist}(f,g) \ge \varepsilon \) for every linear function \(g : \mathbb {Z}_k^n \to \mathbb {Z}_k\). Then for every unit \(j \in (\mathbb {Z}/k\mathbb {Z})^\times \), the \(j\)-twisted lift \(\widetilde{f}_j(x) = \omega _k^{\, j\, f(x)}\) satisfies
where \(\widehat{\widetilde{f}_j}(s)\) denotes the Fourier coefficient of \(\widetilde{f}_j\) at frequency \(s\).
Let \(k\) and \(n\) be natural numbers with \(k \ge 1\). The set \((\mathbb {Z}/k\mathbb {Z})^n\) of \(n\)-tuples with entries in \(\mathbb {Z}/k\mathbb {Z}\) is finite, and its cardinality is
For \(s \in \mathbb {Z}_k^n\), the Fourier character indexed by \(s\) is \(\chi _s : \mathbb {Z}_k^n \to \mathbb {C}\) defined by \(\chi _s(x) = \omega _k^{s \cdot x}\). The \(k^n\) characters \(\{ \chi _s\} _{s \in \mathbb {Z}_k^n}\) form an orthonormal basis of \(L^2(\mathbb {Z}_k^n)\).
Fix integers \(k \ge 1\) and \(n \ge 0\), let \(\omega _k = e^{2\pi i/k}\), and for \(s \in (\mathbb {Z}/k\mathbb {Z})^n\) let \(\chi _s : (\mathbb {Z}/k\mathbb {Z})^n \to \mathbb {C}\) be the character \(\chi _s(x) = \omega _k^{\, s\cdot x}\), where \(s\cdot x = \sum _{i=1}^{n} s_i x_i \in \mathbb {Z}/k\mathbb {Z}\) is the componentwise dot product. Then for all \(x, y \in (\mathbb {Z}/k\mathbb {Z})^n\),
Fix an integer \(k \ge 1\), and for \(s \in \mathbb {Z}_k^n\) let \(\chi _s(x) = \omega _k^{s \cdot x}\) be the Fourier character on \(\mathbb {Z}_k^n\) indexed by \(s\), where \(\omega _k = e^{2\pi i/k}\) and \(s \cdot x = \sum _{i} s_i x_i \in \mathbb {Z}/k\mathbb {Z}\). Then for every \(s, x \in \mathbb {Z}_k^n\),
that is, complex conjugation of a character amounts to negating its index.
Fix an integer \(k \ge 1\) and let \(\omega _k = e^{2\pi i/k}\). For a frequency vector \(s \in \mathbb {Z}_k^n\), let \(\chi _s \colon \mathbb {Z}_k^n \to \mathbb {C}\) be the Fourier character \(\chi _s(x) = \omega _k^{\, s \cdot x}\), where \(s \cdot x = \sum _{i} s_i x_i\) is the dot product in \(\mathbb {Z}/k\mathbb {Z}\). Then for all \(s, t, x \in \mathbb {Z}_k^n\),
so that pointwise multiplication of characters corresponds to addition of their frequency indices.
Fix an integer \(k \ge 1\) and a dimension \(n\), and let \(\chi _s(x) = \omega _k^{s \cdot x}\) denote the Fourier character on \(\mathbb {Z}_k^n\) indexed by \(s \in \mathbb {Z}_k^n\), where \(\omega _k = e^{2\pi i/k}\) and \(s \cdot x = \sum _{i} s_i x_i\) is the dot product modulo \(k\). Then the character indexed by the zero vector is the constant function \(1\): for every \(x \in \mathbb {Z}_k^n\),
Fix an integer \(k \ge 1\) and let \(\omega _k = e^{2\pi i/k}\). For any \(s \in \mathbb {Z}_k^n\), the Fourier character \(\chi _s(x) = \omega _k^{s \cdot x}\) satisfies \(\chi _s(0) = 1\), where \(0\) denotes the zero vector of \(\mathbb {Z}_k^n\).
Fix a modulus \(k \ge 1\) and a dimension \(n\), and let \(\omega _k = e^{2\pi i/k}\). For a frequency \(s \in \mathbb {Z}_k^n\), let \(\chi _s(x) = \omega _k^{s\cdot x}\) be the associated Fourier character on \(\mathbb {Z}_k^n\). If \(s \neq 0\), then the character has vanishing mean:
Fix an integer \(k \ge 1\) and let \(\omega _k = e^{2\pi i/k}\) be the associated primitive \(k\)-th root of unity. For a natural number \(n\) and an index \(s \in (\mathbb {Z}/k\mathbb {Z})^n\), let \(\chi _s \colon (\mathbb {Z}/k\mathbb {Z})^n \to \mathbb {C}\) be the Fourier character \(\chi _s(x) = \omega _k^{\, s \cdot x}\), and write \(\mathbb {E}[f] = k^{-n}\sum _{x} f(x)\) for the uniform average of a complex-valued function \(f\) over \((\mathbb {Z}/k\mathbb {Z})^n\). Then the trivial character \(\chi _0\), indexed by the zero vector, has expectation \(\mathbb {E}[\chi _0] = 1\).
Fix an integer \(k \ge 1\) and a dimension \(n\), and let \(\omega _k = e^{2\pi i / k}\). For a frequency \(s \in \mathbb {Z}_k^n\), write \(\chi _s(x) = \omega _k^{\, s \cdot x}\) for the associated Fourier character on \(\mathbb {Z}_k^n\), where \(s \cdot x = \sum _{i} s_i x_i\) is computed in \(\mathbb {Z}/k\mathbb {Z}\). Then for any two distinct frequencies \(s \ne t\) in \(\mathbb {Z}_k^n\), the Fourier coefficient of the character \(\chi _s\) at \(t\) vanishes:
Fix an integer \(k \ge 1\) and a dimension \(n\), and let \(\chi _s(x) = \omega _k^{s \cdot x}\) denote the Fourier character on \(\mathbb {Z}_k^n\) indexed by a frequency \(s \in \mathbb {Z}_k^n\), where \(\omega _k = e^{2\pi i / k}\). Then for every \(s \in \mathbb {Z}_k^n\), the Fourier coefficient of \(\chi _s\) at its own frequency equals \(1\); equivalently, \(\widehat{\chi _s}(s) = \langle \chi _s, \chi _s \rangle = 1\).
Fix an integer \(k \ge 1\) and let \(f, g : \mathbb {Z}_k^n \to \mathbb {C}\) be complex-valued functions on \(\mathbb {Z}_k^n\), with convolution \((f * g)(x) = \mathbb {E}_y\! \left[f(y)\, g(x - y)\right]\) and Fourier coefficients \(\hat{f}(s) = \langle f, \chi _s\rangle \) taken at any frequency \(s \in \mathbb {Z}_k^n\). Then the Fourier transform carries convolution to pointwise multiplication:
Fix integers \(k \ge 1\) and \(n\), and let \(\omega _k = e^{2\pi i/k}\). For a frequency \(s \in \mathbb {Z}_k^n\), let \(\chi _s(x) = \omega _k^{\, s\cdot x}\) be the corresponding Fourier character, where \(s\cdot x = \sum _{i} s_i x_i \in \mathbb {Z}/k\mathbb {Z}\), and let \(\hat f(s) = \tfrac {1}{k^n}\sum _{x \in \mathbb {Z}_k^n} f(x)\, \overline{\chi _s(x)}\) be the Fourier coefficient of \(f\) at \(s\). Then every function \(f : \mathbb {Z}_k^n \to \mathbb {C}\) is recovered from its Fourier coefficients by
Let \(k \ge 1\) and let \(\omega _k = e^{2\pi i/k}\) be the primitive \(k\)-th root of unity. For each \(a \in \mathbb {Z}/k\mathbb {Z}\), write \(\omega _k^{a}\) for \(\omega _k\) raised to the canonical representative of \(a\) in \(\{ 0,\ldots ,k-1\} \). Then for every \(j \in \mathbb {Z}/k\mathbb {Z}\),
Let \(k \ge 1\) and \(n \ge 0\), and let \(\omega _k = e^{2\pi i/k}\). For \(s, t \in \mathbb {Z}_k^n\) with \(s \ne t\), the Fourier characters \(\chi _s(x) = \omega _k^{s\cdot x}\) and \(\chi _t(x) = \omega _k^{t\cdot x}\) are orthogonal with respect to the Hermitian inner product \(\langle f, g\rangle = \frac{1}{k^n}\sum _{x \in \mathbb {Z}_k^n} f(x)\, \overline{g(x)}\); that is, \(\langle \chi _s, \chi _t\rangle = 0\).
Fix an integer \(k \ge 1\) and a dimension \(n \ge 0\), and let \(\omega _k = e^{2\pi i/k}\). For \(s \in \mathbb {Z}_k^n\), let \(\chi _s(x) = \omega _k^{\, s\cdot x}\) be the Fourier character indexed by \(s\), and equip functions on \(\mathbb {Z}_k^n\) with the Hermitian inner product \(\langle f,g\rangle = \frac{1}{k^n}\sum _{x\in \mathbb {Z}_k^n} f(x)\, \overline{g(x)}\). Then every character has unit norm: for all \(s \in \mathbb {Z}_k^n\),
Let \(k\) be a positive integer and set \(\omega _k = e^{2\pi i / k} \in \mathbb {C}\). Then \(\omega _k\) is a primitive \(k\)-th root of unity: \(\omega _k^k = 1\), and no smaller positive power of \(\omega _k\) equals \(1\), so that \(\omega _k^j = 1\) for a natural number \(j\) holds precisely when \(k\) divides \(j\).
Fix an integer \(k \ge 1\) and let \(\omega _k = e^{2\pi i/k}\) be the primitive \(k\)-th root of unity. For \(s \in \mathbb {Z}_k^n\), let \(\chi _s \colon \mathbb {Z}_k^n \to \mathbb {C}\) be the Fourier character \(\chi _s(x) = \omega _k^{\, s\cdot x}\), where \(s \cdot x = \sum _{i} s_i x_i \in \mathbb {Z}/k\mathbb {Z}\) is the componentwise dot product. Then for every \(s, x \in \mathbb {Z}_k^n\) one has \(\left\lvert \chi _s(x)\right\rvert = 1\).
Fix an integer \(k \ge 1\) and write \(\omega _k = e^{2\pi i/k}\) for the primitive \(k\)-th root of unity. For every residue \(a \in \mathbb {Z}/k\mathbb {Z}\), the complex number \(\omega _k^{a}\), obtained by raising \(\omega _k\) to the canonical representative of \(a\) in \(\{ 0,\dots ,k-1\} \), satisfies \(\lVert \omega _k^{a}\rVert = 1\), so it lies on the unit circle.
Let \(k \ge 1\) and \(n\) be natural numbers, and let \(f : \mathbb {Z}_k^n \to \mathbb {C}\) be a complex-valued function on \(\mathbb {Z}_k^n\). Then the sum of the squared magnitudes of its Fourier coefficients equals its squared \(L^2\) norm:
where \(\hat{f}(s) = \mathbb {E}_{x}\! \left[f(x)\, \overline{\omega _k^{\, s \cdot x}}\right]\) is the Fourier coefficient at frequency \(s\) and \(\| f\| _2^2 = \tfrac {1}{k^n}\sum _{x \in \mathbb {Z}_k^n} |f(x)|^2\).
For \(a : \mathbb {Z}/k\mathbb {Z}\), \(\texttt{ZkFourier.toOmega}\ a = \omega _k^{a}\), where \(a\) is lifted to its canonical representative in \(\{ 0, \ldots , k-1\} \subset \mathbb {N}\). This is the group homomorphism \((\mathbb {Z}/k\mathbb {Z}, +) \to (\mathbb {C}^\times , \cdot )\).
Fix an integer \(k \ge 1\) and let \(\omega _k = e^{2\pi i/k}\) be the primitive \(k\)-th root of unity. For \(a \in \mathbb {Z}/k\mathbb {Z}\) write \(\chi (a) = \omega _k^{a}\), where the exponent is the canonical representative of \(a\) in \(\{ 0, 1, \ldots , k-1\} \). Then for all \(a, b \in \mathbb {Z}/k\mathbb {Z}\),
so \(\chi \) is a homomorphism from the additive group \(\mathbb {Z}/k\mathbb {Z}\) to the multiplicative group \(\mathbb {C}^{\times }\).
Fix a positive integer \(k\) and let \(\omega _k = e^{2\pi i / k}\). For \(b \in \mathbb {Z}/k\mathbb {Z}\), write \(\lfloor b\rfloor \in \{ 0,\ldots ,k-1\} \) for its canonical representative and set \(\chi (b) = \omega _k^{\lfloor b\rfloor }\). Then for all \(j, a \in \mathbb {Z}/k\mathbb {Z}\),
Let \(k \ge 1\) and let \(\omega _k = e^{2\pi i/k}\) be the primitive \(k\)-th root of unity, and for \(a \in \mathbb {Z}/k\mathbb {Z}\) write \(\omega _k^{a}\) for the value \(\omega _k\) raised to the canonical representative of \(a\) in \(\{ 0,\dots ,k-1\} \). Then for every \(a \in \mathbb {Z}/k\mathbb {Z}\),
that is, replacing \(a\) by its additive inverse conjugates the corresponding character value.
Let \(k \ge 1\) and write \(\omega _k = e^{2\pi i/k}\) for the primitive \(k\)-th root of unity, and for \(a \in \mathbb {Z}/k\mathbb {Z}\) let its image under the homomorphism \((\mathbb {Z}/k\mathbb {Z}, +) \to (\mathbb {C}^\times , \cdot )\) be \(\omega _k^{a}\). Then the additive identity \(0 \in \mathbb {Z}/k\mathbb {Z}\) maps to \(1\); that is, \(\omega _k^{0} = 1\).
Fix an integer \(k \ge 1\) and a dimension \(n\), and let \(s, x, y \in (\mathbb {Z}/k\mathbb {Z})^n\). Writing \(s \cdot x = \sum _{i=0}^{n-1} s_i x_i \in \mathbb {Z}/k\mathbb {Z}\) for the componentwise dot product, the dot product is additive in its right argument:
where \(x + y\) denotes the componentwise sum.
Fix an integer \(k \ge 1\) and a dimension \(n\), and work in the module \(\mathbb {Z}_k^n = (\mathbb {Z}/k\mathbb {Z})^n\) with dot product \(s \cdot x = \sum _{i=0}^{n-1} s_i x_i \in \mathbb {Z}/k\mathbb {Z}\). Then for all \(s, x \in \mathbb {Z}_k^n\),
For any three vectors \(s, t, x \in \mathbb {Z}_k^n\), the dot product is additive in its left argument under subtraction:
where the subtraction on the left is taken componentwise and both sides are elements of \(\mathbb {Z}/k\mathbb {Z}\).
Fix an integer \(k \ge 1\) and a length \(n \ge 0\), and let \(s \in (\mathbb {Z}/k\mathbb {Z})^n\). Writing \(0\) for the zero vector of \((\mathbb {Z}/k\mathbb {Z})^n\), the dot product of \(s\) with \(0\) equals \(0\) in \(\mathbb {Z}/k\mathbb {Z}\); that is, \(\sum _{i=0}^{n-1} s_i \cdot 0 = 0\).
\(\texttt{ZkFourier.ZkVec}\ k\ n\) is the \(n\)-dimensional vector space over \(\mathbb {Z}/k\mathbb {Z}\), defined as the function type \(\mathrm{Fin}\, n \to \mathbb {Z}/k\mathbb {Z}\). It carries the pointwise additive group structure of \((\mathbb {Z}/k\mathbb {Z})^n\).