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CSL853: Complexity Theory CSL853: Complexity Theory Ragesh Jaiswal, CSE, IIT Delhi Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Complexity Diagram Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Complexity Diagram Our current view of complexity classes Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation BPP examples: Perfect matching Definition (MATCHING) MATCHING = fhG = (V1; V2; E)i : G is a bipartite graph that has a perfect matchingg. Theorem MATCHING 2 BPP. Theorem MATCHING 2 randomNC. Proof idea Let A be a matrix such that Ai;j contains the variable xij if there is an edge from vertex i on the left to vertex j on the right, else Ai;j has 0. Claim 1: G has a perfect matching if and only if det(A) is not identically zero polynomial. Claim 2: Language fhM; ki : M is a matrix with integer entries with determinant kg is in NC. Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation BPP examples: Perfect matching Definition (MATCHING) MATCHING = fhG = (V1; V2; E)i : G is a bipartite graph that has a perfect matchingg. Theorem MATCHING 2 BPP. Theorem MATCHING 2 randomNC. Open Problem ? MATCHING 2 NC. Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation RP; coRP; ZPP BPP captures probabilistic algorithms with two-sided error. In certain applications, we might need one-sided error. Definition (RTIME) RTIME(T (n)) contains every language L for which there is a probabilistic TM M running in T (n) times such that x 2 L ) Pr[M(x) = 1] ≥ 2=3 x 2= L ) Pr[M(x) = 1] = 0: Definition (RP) c RP = [c>0RTIME(n ). Definition (coRP) coRP = fL : L 2 RPg. Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation RP; coRP; ZPP BPP captures probabilistic algorithms with two-sided error. In certain applications, we might need one-sided error. In certain applications, we might need zero-sided error. For any PTM M on input x, the running time of the machine on x denoted by TM;x is a random variable. A PTM M is said to have expected running time T (n) if the ∗ expectation E[TM;x ] is at most T (jxj) for any x 2 f0; 1g . Definition (ZTIME) The class ZTIME(T (n)) contains all the languages L for which there is a PTM M that runs in expected time O(T (n)) such that for every input x, whenever M halts on x, the output M(x) it produces is exactly L(x). Definition (ZPP) c ZPP = [c>0ZTIME(n ). Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation RP; coRP; ZPP Definition (ZTIME) The class ZTIME(T (n)) contains all the languages L for which there is a PTM M that runs in expected time O(T (n)) such that for every input x, whenever M halts on x, the output M(x) it produces is exactly L(x). Definition (ZPP) c ZPP = [c>0ZTIME(n ). Theorem ZPP = RP \ coRP. Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation BPP; RP; coRP; ZPP Here is the relationship between the randomised classes. How does this new structure fit into our known complexity diagram? Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation BPP; RP; coRP; ZPP Claim 1:P ⊆ ZPP. Claim 2: RP ⊆ NP. Claim 3: coRP ⊆ coNP. How is BPP related to NP? Is it possible that NP ⊆ BPP? Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation BPP; RP; coRP; ZPP Claim 1:P ⊆ BPP. Claim 2: RP ⊆ NP. Claim 3: coRP ⊆ coNP. How is BPP related to NP? Theorem (Adleman's Theorem, 1978) BPP ⊆ P=poly. Is it possible that NP ⊆ BPP? This is unlikely. Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation BPP; RP; coRP; ZPP Theorem (Adleman's Theorem, Adleman, 1978) BPP ⊆ P=poly. Proof. Claim 1: For any L 2 BPP, there is a M and a polynomial p such that M runs in polynomial time such that 1 Pr p(jxj) [M(x; r) = L(x)] > 1 − : r2R f0;1g 2jxj+1 Theorem (Chernoff Bounds) Let X1; X2; :::; Xn be mutually independent random variables over f0; 1g Pn and let µ = i=1 E[Xi ]. Then for every c > 0, " n # X −µ·min (c2=4;c=2) Pr j Xi − µj ≥ cµ ≤ 2 · e i=1 Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation BPP; RP; coRP; ZPP Theorem (Adleman's Theorem, Adleman, 1978) BPP ⊆ P=poly. Proof. Claim 1: For any L 2 BPP, there is a M and a polynomial p such that M runs in polynomial time such that 1 Pr p(jxj) [M(x; r) = L(x)] > 1 − : r2R f0;1g 2jxj+1 Consider inputs of size n. Let m = p(n). Let r 2 f0; 1gm be called \bad" if there is an input x 2 f0; 1gn such that M(x; r) 6= L(x), otherwise it is called \good". Claim 2: For every n, there is a string r 2 f0; 1gm that is good for all inputs of length n. Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation BPP; RP; coRP; ZPP Theorem (Adleman's Theorem, Adleman, 1978) BPP ⊆ P=poly. Proof. Claim 1: For any L 2 BPP, there is a M and a polynomial p such that M runs in polynomial time such that 1 Pr p(jxj) [M(x; r) = L(x)] > 1 − : r2R f0;1g 2jxj+1 Consider inputs of size n. Let m = p(n). Let r 2 f0; 1gm be called \bad" if there is an input x 2 f0; 1gn such that M(x; r) 6= L(x), otherwise it is called \good". Claim 2: For every n, there is a string r 2 f0; 1gp(n) that is good for all inputs of length n. This shows that BPP ⊆ P=poly since such strings may be used as polynomial size \advice". Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation P P BPP ⊆ Σ2 \ Π2 Theorem (Sipser-G`acs-Lautemanntheorem, 1983) P P BPP ⊆ Σ2 \ Π2 . Proof. P Claim 1: It is sufficient to show that BPP ⊆ Σ2 . Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation P P BPP ⊆ Σ2 \ Π2 Theorem (Sipser-G`acs-Lautemanntheorem, 1983) P P BPP ⊆ Σ2 \ Π2 . Proof. P Claim 1: It is sufficient to show that BPP ⊆ Σ2 . Claim 2: For any L 2 BPP, there is a TM M that on inputs of length n uses m = poly(n) random bits such that −n x 2 L ) Prr2f0;1gm [M(x; r) accepts] ≥ 1 − 2 −n x 2= L ) Prr2f0;1gm [M(x; r) accepts] ≤ 2 : Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation P P BPP ⊆ Σ2 \ Π2 Theorem (Sipser-G`acs-Lautemanntheorem, 1983) P P BPP ⊆ Σ2 \ Π2 . Proof. P Claim 1: It is sufficient to show that BPP ⊆ Σ2 . Claim 2: For any L 2 BPP, there is a TM M that on inputs of length n uses m = poly(n) random bits such that −n x 2 L ) Prr2f0;1gm [M(x; r) accepts] ≥ 1 − 2 −n x 2= L ) Prr2f0;1gm [M(x; r) accepts] ≤ 2 : For a set S ⊆ f0; 1gm and string u 2 f0; 1gm, let S + u = fx + u : x 2 Sg, where + denotes the bitwise XOR operation. Let k = dm=ne + 1. Claim 3: For every set S ⊆ f0; 1gm with jSj ≤ 2m−n and every k m k m vectors u1; :::; uk 2 f0; 1g , [i=1(S + ui ) 6= f0; 1g . Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation P P BPP ⊆ Σ2 \ Π2 Theorem (Sipser-G`acs-Lautemanntheorem, 1983) P P BPP ⊆ Σ2 \ Π2 . Proof. P Claim 1: It is sufficient to show that BPP ⊆ Σ2 . Claim 2: For any L 2 BPP, there is a TM M that on inputs of length n uses m = poly(n) random bits such that −n x 2 L ) Prr2f0;1gm [M(x; r) accepts] ≥ 1 − 2 −n x 2= L ) Prr2f0;1gm [M(x; r) accepts] ≤ 2 : For a set S ⊆ f0; 1gm and string u 2 f0; 1gm, let S + u = fx + u : x 2 Sg, where + denotes the bitwise XOR operation. Let k = dm=ne + 1. Claim 3: For every set S ⊆ f0; 1gm with jSj ≤ 2m−n and every k m k m vectors u1; :::; uk 2 f0; 1g , [i=1(S + ui ) 6= f0; 1g . Claim 4: For every set S ⊆ f0; 1gm with jSj ≥ (1 − 2−n)2m, there m k m exists u1; :::; uk 2 f0; 1g such that [i=1(S + ui ) = f0; 1g . Ragesh Jaiswal, CSE, IIT Delhi CSL853: Complexity Theory Randomized Computation P P BPP ⊆ Σ2 \ Π2 Theorem (Sipser-G`acs-Lautemanntheorem, 1983) P P BPP ⊆ Σ2 \ Π2 . Proof. P Claim 1: It is sufficient to show that BPP ⊆ Σ2 . Claim 2: For any L 2 BPP, there is a TM M that on inputs of length n uses m = poly(n) random bits such that −n x 2 L ) Prr2f0;1gm [M(x; r) accepts] ≥ 1 − 2 −n x 2= L ) Prr2f0;1gm [M(x; r) accepts] ≤ 2 : For a set S ⊆ f0; 1gm and string u 2 f0; 1gm, let S + u = fx + u : x 2 Sg, where + denotes the bitwise XOR operation. Let k = dm=ne + 1. Claim 3: For every set S ⊆ f0; 1gm with jSj ≤ 2m−n and every k m k m vectors u1; :::; uk 2 f0; 1g , [i=1(S + ui ) 6= f0; 1g . Claim 4: For every set S ⊆ f0; 1gm with jSj ≥ (1 − 2−n)2m, there m k m exists u1; :::; uk 2 f0; 1g such that [i=1(S + ui ) = f0; 1g .
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