# 6.045: Automata, Computability, and Complexity (GITCS) Class 17 Nancy Lynch

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1 6.045: Automata, Computability, and Complexity (GITCS) Class 17 Nancy Lynch

2 Today Probabilistic Turing Machines and Probabilistic Time Complexity Classes Now add a new capability to standard TMs: random choice of moves. Gives rise to new complexity classes: BPP and RP Topics: Probabilistic polynomial-time TMs, BPP and RP Amplification lemmas Example 1: Primality testing Example 2: Branching-program equivalence Relationships between classes Reading: Sipser Section 10.2

3 Probabilistic Polynomial-Time Turing Machines, BPP and RP

4 Probabilistic Polynomial-Time TM New kind of NTM, in which each nondeterministic step is a coin flip: has exactly 2 next moves, to each of which we assign probability ½. Example: To each maximal branch, we assign a probability: ½ ½ ½ Computation on input w number of coin flips on the branch Has accept and reject states, as for NTMs. Now we can talk about probability of acceptance or rejection, on input w. 1/16 1/16 1/4 1/4 1/8 1/8 1/8

5 Probabilistic Poly-Time TMs Probability of acceptance = Σ b an accepting branch Pr(b) Probability of rejection = Σ b a rejecting branch Pr(b) Example: Add accept/reject information Probability of acceptance = 1/16 + 1/8 + 1/4 + 1/8 + 1/4 = 13/16 Probability of rejection = 1/16 + 1/8 = 3/16 We consider TMs that halt (either accept or reject) on every branch-- -deciders. So the two probabilities total 1. Computation on input w 1/4 1/4 1/8 1/8 1/8 1/16 1/16 Acc Acc Acc Acc Rej Acc Rej

6 Probabilistic Poly-Time TMs Time complexity: Worst case over all branches, as usual. Q: What good are probabilistic TMs? Random choices can help solve some problems efficiently. Good for getting estimates---arbitrarily accurate, based on the number of choices. Example: Monte Carlo estimation of areas E.g, integral of a function f. Repeatedly choose a random point (x,y) in the rectangle. Compare y with f(x). Fraction of trials in which y f(x) can be used to estimate the integral of f. f

7 Probabilistic Poly-Time TMs Random choices can help solve some problems efficiently. We ll see 2 languages that have efficient probabilistic estimation algorithms. Q: What does it mean to estimate a language? Each w is either in the language or not; what does it mean to approximate a binary decision? Possible answer: For most inputs w, we always get the right answer, on all branches of the probabilistic computation tree. Or: For most w, we get the right answer with high probability. Better answer: For every input w, we get the right answer with high probability.

8 Probabilistic Poly-Time TMs Better answer: For every input w, we get the right answer with high probability. Definition: A probabilistic TM decider M decides language L with error probability ε if w L implies that Pr[ M accepts w ] 1 - ε, and w L implies that Pr[ M rejects w ] 1 - ε. Definition: Language L is in BPP (Bounded-error Probabilistic Polynomial time) if there is a probabilistic polynomial-time TM that decides L with error probability 1/3. Q: What s so special about 1/3? Nothing. We would get an equivalent definition (same language class) if we chose ε to be any value with 0 < ε < ½. We ll see this soon---amplification Theorem

9 Probabilistic Poly-Time TMs Another class, RP, where the error is 1-sided: Definition: Language L is in RP (Random Polynomial time) if there is a a probabilistic polynomial-time TM that decides L, where: w L implies that Pr[ M accepts w ] 1/2, and w L implies that Pr[ M rejects w ] = 1. Thus, absolutely guaranteed to be correct for words not in L---always rejects them. But, might be incorrect for words in L---might mistakenly reject these, in fact, with probability up to ½. We can improve the ½ to any larger constant < 1, using another Amplification Theorem.

10 RP Definition: Language L is in RP (Random Polynomial time) if there is a a probabilistic polynomial-time TM that decides L, where: w L implies that Pr[ M accepts w ] 1/2, and w L implies that Pr[ M rejects w ] = 1. Always correct for words not in L. Might be incorrect for words in L---can reject these with probability up to ½. Compare with nondeterministic TM acceptance: w L implies that there is some accepting path, and w L implies that there is no accepting path.

11 Amplification Lemmas

12 Amplification Lemmas Lemma: Suppose that M is a PPT-TM that decides L with error probability ε, where 0 ε < ½. Then for any ε, 0 ε < ½, there exists M, another PPT- TM, that decides L with error probability ε. Proof idea: M simulates M many times and takes the majority value for the decision. Why does this improve the probability of getting the right answer? E.g., suppose ε = 1/3; then each trial gives the right answer at least 2/3 of the time (with 2/3 probability). If we repeat the experiment many times, then with very high probability, we ll get the right answer a majority of the times. How many times? Depends on ε and ε.

13 Amplification Lemmas Lemma: Suppose that M is a PPT-TM that decides L with error probability ε, where 0 ε < ½. Then for any ε, 0 ε < ½, there exists M, another PPT- TM, that decides L with error probability ε. Proof idea: M simulates M many times, takes the majority value. E.g., suppose ε = 1/3; then each trial gives the right answer at least 2/3 of the time (with 2/3 probability). If we repeat the experiment many times, then with very high probability, we ll get the right answer a majority of the times. How many times? Depends on ε and ε. 2k, where ( 4ε (1- ε) ) k ε, suffices. In other words k (log 2 ε ) / (log 2 (4ε (1- ε))). See book for calculations.

14 Characterization of BPP Theorem: L BPP if and only for, for some ε, 0 ε < ½, there is a PPT-TM that decides L with error probability ε. Proof: If L BPP, then there is some PPT-TM that decides L with error probability ε = 1/3, which suffices. If for some ε, a PPT-TM decides L with error probability ε, then by the Lemma, there is a PPT-TM that decides L with error probability 1/3; this means that L BPP.

15 Amplification Lemmas For RP, the situation is a little different: If w L, then Pr[ M accepts w ] could be equal to ½. So after many trials, the majority would be just as likely to be correct or incorrect. But this isn t useless, because when w L, the machine always answers correctly. Lemma: Suppose M is a PPT-TM that decides L, 0 ε < 1, and w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1. Then for any ε, 0 ε < 1, there exists M, another PPT-TM, that decides L with: w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1.

16 Amplification Lemmas Lemma: Suppose M is a PPT-TM that decides L, 0 ε< 1, w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1. Then for any ε, 0 ε < 1, there exists M, another PPT-TM, that decides L with: w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1. Proof idea: M : On input w: Run k independent trials of M on w. If any accept, then accept; else reject. Here, choose k such that ε k ε. If w L then all trials reject, so M rejects, as needed. If w L then each trial accepts with probability 1 - ε, so Prob(at least one of the k trials accepts) = 1 Prob(all k reject) 1 - ε k 1 - ε.

17 Characterization of RP Lemma: Suppose M is a PPT-TM that decides L, 0 ε< 1, w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1. Then for any ε, 0 ε < 1, there exists M, another PPT- TM, that decides L with: w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1. Theorem: L RP iff for some ε, 0 ε< 1, there is a PPT- TM that decides L with: w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1.

18 RP vs. BPP Lemma: Suppose M is a PPT-TM that decides L, 0 ε< 1, w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1. Then for any ε, 0 ε < 1, there exists M, another PPT-TM, that decides L with: w L implies Pr[ M accepts w ] 1 - ε. w L implies Pr[ M rejects w ] = 1. Theorem: RP BPP. Proof: Given A RP, get (by def. of RP) a PPT-TM M with: w L implies Pr[ M accepts w ] ½. w L implies Pr[ M rejects w ] = 1. By Lemma, get another PPT-TM for A, with: w L implies Pr[ M accepts w ] 2/3. w L implies Pr[ M rejects w ] = 1. Implies A BPP, by definition of BPP.

19 RP, co-rp, and BPP Definition: corp = { L L c RP } corp contains the languages L that can be decided by a PPT-TM that is always correct for w L and has error probability at most ½ for w L. That is, L is in corp if there is a PPT-TM that decides L, where: w L implies that Pr[ M accepts w ] = 1, and w L implies that Pr[ M rejects w ] 1/2. Theorem: corp BPP. So we have: BPP RP corp

20 Example 1: Primality Testing

21 Primality Testing PRIMES = { <n> n is a natural number > 1 and n cannot be factored as q r, where 1 < q, r < n } COMPOSITES = { <n> n > 1 and n can be factored } We will show an algorithm demonstrating that PRIMES corp. So COMPOSITES RP, and both BPP. BPP COMPOSITES RP corp PRIMES This is not exciting, because it is now known that both are in P. [Agrawal, Kayal, Saxema 2002] But their poly-time algorithm is hard, whereas the probabilistic algorithm is easy. And anyway, this illustrates some nice probabilistic methods.

22 Primality Testing PRIMES = { <n> n is a natural number > 1 and n cannot be factored as q r, where 1 < q, r < n } COMPOSITES = { <n> n > 1 and n can be factored } BPP COMPOSITES RP corp PRIMES Note: Deciding whether n is prime/composite isn t the same as factoring. Factoring seems to be a much harder problem; it s at the heart of modern cryptography.

23 Primality Testing PRIMES = { <n> n is a natural number > 1 and n cannot be factored as q r, where 1 < q, r < n } Show PRIMES corp. Design PPT-TM (algorithm) M for PRIMES that satisfies: n PRIMES Pr[ M accepts n] = 1. n PRIMES Pr[ M accepts n] 2 -k. Here, k depends on the number of trials M makes. M always accepts primes, and almost always correctly identifies composites. Algorithm rests on some number-theoretic facts about primes (just state them here):

24 Fermat s Little Theorem PRIMES = { <n> n is a natural number > 1 and n cannot be factored as q r, where 1 < q, r < n } Design PPT-TM (algorithm) M for PRIMES that satisfies: n PRIMES Pr[ M accepts n] = 1. n PRIMES Pr[ M accepts n] 2 -k. Fact 1: Fermat s Little Theorem: If n is prime and a Z n + then a n-1 1 mod n. Example: n = 5, Z n+ = {1,2,3,4}. a = 1: = 1 4 = 1 1 mod 5. a = 2: = 2 4 = 16 1 mod 5. a = 3: = 3 4 = 81 1 mod 5. a = 4: = 4 4 = mod 5. Integers mod n except for 0, that is, {1,2,,n-1}

25 Fermat s test Design PPT-TM (algorithm) M for PRIMES that satisfies: n PRIMES Pr[ M accepts n] = 1. n PRIMES Pr[ M accepts n] 2 -k. Fermat: If n is prime and a Z n+ then a n-1 1 mod n. We can use this fact to identify some composites without factoring them: Example: n = 8, a = = mod 8, not 1 mod 8. So 8 is composite. Algorithm attempt 1: On input n: Choose a number a randomly from Z n+ = { 1,,n-1 }. If a n-1 1 mod n then accept (passes Fermat test). Else reject (known not to be prime).

26 Algorithm attempt 1 Design PPT-TM (algorithm) M for PRIMES that satisfies: n PRIMES Pr[ M accepts n] = 1. n PRIMES Pr[ M accepts n] 2 -k. Fermat: If n is prime and a Z n+ then a n-1 1 mod n. First try: On input n: Choose number a randomly from Z n+ = { 1,,n-1 }. If a n-1 1 mod n then accept (passes Fermat test). Else reject (known not to be prime). This guarantees: n PRIMES Pr[ M accepts n] = 1. n PRIMES?? Don t know. It could pass the test, and be accepted erroneously. The problem isn t helped by repeating the test many times, for many values of a---because there are some non-prime n s that pass the test for all values of a.

27 Carmichael numbers Fermat: If n is prime and a Z n+ then a n-1 1 mod n. On input n: Choose a randomly from Z n+ = { 1,,n-1 }. If a n-1 1 mod n then accept (passes Fermat test). Else reject (known not to be prime). Carmichael numbers: Non-primes that pass all Fermat tests, for all values of a. Fact 2: Any non-carmichael composite number fails at least half of all Fermat tests (for at least half of all values of a). So for any non-carmichael composite, the algorithm correctly identifies it as composite, with probability ½. So, we can repeat k times to get more assurance. Guarantees: n PRIMES Pr[ M accepts n] = 1. n a non-carmichael composite number Pr[ M accepts n] 2 -k. n a Carmichael composite number Pr[ M accepts n ] = 1 (wrong)

28 Carmichael numbers Fermat: If n is prime and a Z n+ then a n-1 1 mod n. On input n: Choose a randomly from Z n+ = { 1,,n-1 }. If a n-1 1 mod n then accept (passes Fermat test). Else reject (known not to be prime). Carmichael numbers: Non-primes that pass all Fermat tests. Algorithm guarantees: n PRIMES Pr[ M accepts n] = 1. n a non-carmichael composite number Pr[ M accepts n] 2 -k. n a Carmichael composite number Pr[ M accepts n] = 1. We must do something about the Carmichael numbers. Use another test, based on: Fact 3: For every Carmichael composite n, there is some b 1, -1 such that b 2 1 mod n (that is, 1 has a nontrivial square root, mod n). No prime has such a square root.

29 Primality-testing algorithm Fact 3: For every Carmichael composite n, there is some b 1, -1 such that b 2 1 mod n. No prime has such a square root. Primality-testing algorithm: On input n: If n = 1 or n is even: Give the obvious answer (easy). If n is odd and > 1: Choose a randomly from Z n+. (Fermat test) If a n-1 is not congruent to 1 mod n then reject. (Carmichael test) Write n 1 = 2 h s, where s is odd (factor out twos). Consider successive squares, a s, a 2s, a 4s, a 8s..., a 2^h s = a n-1. If all terms are 1 mod n, then accept. If not, then find the last one that isn t congruent to 1. If it s -1 mod n then accept else reject.

30 Primality-testing algorithm If n is odd and > 1: Choose a randomly from Z n+. (Fermat test) If a n-1 is not congruent to 1 mod n then reject. (Carmichael test) Write n 1 = 2 h s, where s is odd. Consider successive squares, a s, a 2s, a 4s, a 8s..., a 2^h s = a n-1. If all terms are 1 mod n, then accept. If not, then find the last one that isn t congruent to 1. If it s -1 mod n then accept else reject. Theorem: This algorithm satisfies: n PRIMES Pr[ accepts n] = 1. n PRIMES Pr[ accepts n] ½. By repeating it k times, we get: n PRIMES Pr[ accepts n] (½) k.

31 Primality-testing algorithm If n is odd and > 1: Choose a randomly from Z n+. (Fermat test) If a n-1 is not congruent to 1 mod n then reject. (Carmichael test) Write n 1 = 2 h s, where s is odd. Consider successive squares, a s, a 2s, a 4s, a 8s..., a 2^h s = a n-1. If all terms are 1 mod n, then accept. If not, then find the last one that isn t congruent to 1. If it s -1 mod n then accept else reject. Theorem: This algorithm satisfies: n PRIMES Pr[ accepts n] = 1. n PRIMES Pr[ accepts n] ½. Proof: Suppose n is odd and > 1.

32 Proof If n is odd and > 1: Choose a randomly from Z n+. (Fermat test) If a n-1 is not congruent to 1 mod n then reject. (Carmichael test) Write n 1 = 2 h s, where s is odd. Consider successive squares, a s, a 2s, a 4s, a 8s..., a 2^h s = a n-1. If all terms are 1 mod n, then accept. If not, then find the last one that isn t congruent to 1. If it s -1 mod n then accept else reject. Proof that n PRIMES Pr[accepts n] = 1. Show that, if the algorithm rejects, then n must be composite. Reject because of Fermat: Then not prime, by Fact 1 (primes pass). Reject because of Carmichael: Then 1 has a nontrivial square root b, mod n, so n isn t prime, by Fact 3. Let b be the last term in the sequence that isn t congruent to 1 mod n. b 2 is the next one, and is 1 mod n, so b is a square root of 1, mod n.

33 Proof If n is odd and > 1: Choose a randomly from Z n+. (Fermat test) If a n-1 is not congruent to 1 mod n then reject. (Carmichael test) Write n 1 = 2 h s, where s is odd. Consider successive squares, a s, a 2s, a 4s, a 8s..., a 2^h s = a n-1. If all terms are 1 mod n, then accept. If not, then find the last one that isn t congruent to 1. If it s -1 mod n then accept else reject. Proof that n PRIMES Pr[accepts n] ½. Suppose n is a composite. If n is not a Carmichael number, then at least half of the possible choices of a fail the Fermat test (by Fact 2). If n is a Carmichael number, then Fact 3 says that some b fails the Carmichael test (is a nontrivial square root). Actually, when we generate b using a as above, at least half of the possible choices of a generate bs that fail the Carmichael test. Why: Technical argument, in Sipser, p

34 Primality-testing algorithm So we have proved: Theorem: This algorithm satisfies: n PRIMES Pr[ accepts n] = 1. n PRIMES Pr[ accepts n] ½. This implies: Theorem: PRIMES corp. Repeating k times, or using an amplification lemma, we get: n PRIMES Pr[ accepts n] = 1. n PRIMES Pr[ accepts n] (½) k. Thus, the algorithm might sometimes make mistakes and classify a composite as a prime, but the probability of doing this can be made arbitrarily low. Corollary: COMPOSITES RP.

35 Primality-testing algorithm Theorem: PRIMES corp. Corollary: COMPOSITES RP. Corollary: Both PRIMES and COMPOSITES BPP. BPP COMPOSITES RP corp PRIMES

36 Example 2: Branching-Program Equivalence

37 Branching Programs Branching program: A variant of a decision tree. Can be a DAG, not just a tree: Describes a Boolean function of a set { x 1, x 2, x 3, } of Boolean variables. Restriction: Each variable appears at most once on each path. Example: x 1 x 2 x 3 result x x 2 x x 3 x

38 Branching Programs Branching program representation for Boolean functions is used by system modeling and analysis tools, for systems in which the state can be represented using just Boolean variables. Programs called Binary Decision Diagrams (BDDs). Analyzing a model involves exploring all the states, which in turn involves exploring all the paths in the diagram. Choosing the right order of evaluating the variables can make a big difference in cost (running time). Q: Given two branching programs, B 1 and B 2, do they compute the same Boolean function? That is, do the same values for all the variables always lead to the same result in both programs?

39 Branching-Program Equivalence Q: Given two branching programs, B 1 and B 2, do they compute the same Boolean function? Express as a language problem: EQ BP = { < B 1, B 2 > B 1 and B 2 are BPs that compute the same Boolean function }. Theorem: EQ BP is in corp BPP. Note: Need the restriction that a variable appears at most once on each path. Otherwise, the problem is conpcomplete. Proof idea: Pick random values for x 1, x 2, and see if they lead to the same answer in B 1 and B 2. If so, accept; if not, reject. Repeat several times for extra assurance.

40 Branching-Program Equivalence EQ BP = { < B 1, B 2 > B 1 and B 2 are BPs that compute the same Boolean function } Theorem: EQ BP is in corp BPP. Proof idea: Pick random values for x 1, x 2, and see if they lead to the same answer in B 1 and B 2. If so, accept; if not, reject. Repeat several times for extra assurance. This is not quite good enough: Some inequivalent BPs differ on only one assignment to the vars. Unlikely that the algorithm would guess this assignment. Better proof idea: Consider the same BPs but now pretend the domain of values for the variables is Z p, the integers mod p, for a large prime p, rather than just {0,1}. This will let us make more distinctions, making it less likely that we would think B 1 and B 2 are equivalent if they aren t.

41 Branching-Program Equivalence EQ BP = { < B 1, B 2 > B 1 and B 2 are BPs that compute the same Boolean function } Theorem: EQ BP is in corp BPP. Proof idea: Pick random values for x 1, x 2, and see if they lead to the same answer in B 1 and B 2. If so, accept; if not, reject. Repeat several times for extra assurance. Better proof idea: Pretend that the domain of values for the variables is Z p, the integers mod p, for a large prime p, rather than just {0,1}. This lets us make more distinctions, making it less likely that we would think B 1 and B 2 are equivalent if they aren t. But how do we apply the programs to integers mod p? By associating a multi-variable polynomial with each program:

42 Associating a polynomial with a BP Associate a polynomial with each node in the BP, and use the poly associated with the 1-result node as the poly for the entire BP. 1 0 x x 1 x 2 x x 1 (1-x 1 )(1-x 2 ) (1-x 1 ) (1-x 2 ) (1-x 3 ) + (1- x 1 ) x 2 + x 1 (1-x 3 ) (1- x 2 ) x x x 1 (1-x 3 ) x 1 (1-x 3 ) x 2 + x 1 x 3 + (1-x 1 ) (1-x 2 ) x 3 The polynomial associated with the program

43 Labeling rules Top node: Label with polynomial 1. Non-top node: Label with sum of polys, one for each incoming edge: Edge labeled with 1, from x, labeled with p, contributes p x. Edge labeled with 0, from x, labeled with p, contributes p (1-x). 1 0 x x 1 x 2 x x 1 (1-x 1 )(1-x 2 ) (1-x 1 ) (1-x 2 ) (1-x 3 ) + (1- x 1 ) x 2 + x 1 (1-x 3 ) (1- x 2 ) x x x 1 (1-x 3 ) x 1 (1-x 3 ) x 2 + x 1 x 3 + (1-x 1 ) (1-x 2 ) x 3 The polynomial associated with the program

44 Labeling rules Top node: Label with polynomial 1. Non-top node: Label with sum of polys, one for each incoming edge: Edge labeled with 1, from x labeled with p, contributes p x. Edge labeled with 0, from x labeled with p, contributes p (1-x). x p x p 1 0 p x p (1-x)

45 Associating a polynomial with a BP What do these polynomials mean for Boolean values? For any particular assignment of { 0, 1 } to the variables, each polynomial at each node evaluates to either 0 or 1 (because of their special form). The polynomials on the path followed by that assignment all evaluate to 1, and all others evaluate to 0. The polynomial associated with the entire program evaluates to 1 exactly for the assignments that lead there = those that are assigned value 1 by the program. Example: Above. The assignments leading to result 1 are: Which are exactly the assignments for which the program s polynomial evaluates to 1. x 1 (1-x 3 ) x 2 + x 1 x 3 + (1-x 1 ) (1-x 2 ) x 3 x 1 x 2 x

46 Branching-Program Equivalence Now consider Z p, integers mod p, for a large prime p (much bigger than the number of variables). Equivalence algorithm: On input < B 1, B 2 >, where both programs use m variables: Choose elements a 1, a 2,,a m from Z p at random. Evaluate the polynomials p 1 associated with B 1 and p 2 associated with B 2 for x 1 = a 1, x 2 = a 2,,x m = a m. Evaluate them node-by-node, without actually constructing all the polynomials for both programs. Do this in polynomial time in the size of < B 1, B 2 >, LTTR. If the results are equal (mod p) then accept; else reject. Theorem: The equivalence algorithm guarantees: If B 1 and B 2 are equivalent BPs (for Boolean values) then Pr[ algorithm accepts n] = 1. If B 1 and B 2 are not equivalent, then Pr[ algorithm rejects n] 2/3.

47 Branching-Program Equivalence Equivalence algorithm: On input < B 1, B 2 >: Choose elements a 1, a 2,,a m from Z p at random. Evaluate the polynomials p 1 associated with B 1 and p 2 associated with B 2 for x 1 = a 1, x 2 = a 2,,x m = a m. If the results are equal (mod p) then accept; else reject. Theorem: The equivalence algorithm guarantees: If B 1 and B 2 are equivalent BPs then Pr[ accepts n] = 1. If B 1 and B 2 are not equivalent, then Pr[ rejects n] 2/3. Proof idea: (See Sipser, p. 379) If B 1 and B 2 are equivalent BPs (for Boolean values), then p 1 and p 2 are equivalent polynomials over Z p, so always accepts. If B 1 and B 2 are not equivalent (for Boolean values), then at least 2/3 of the possible sets of choices from Z p yield different values, so Pr[ rejects n] 2/3. Corollary: EQ BP corp BPP.

48 Relationships Between Complexity Classes

49 Relationships between complexity classes We know: BPP RP P corp Also recall: NP conp P From the definitions, RP NP and corp conp. So we have:

50 Relationships between classes So we have: NP RP corp conp P Q:Where does BPP fit in?

51 Relationships between classes Where does BPP fit? NP conp BPP? NP BPP = P? Something in between? RP corp Many people believe P BPP = RP = corp = P, that is, that randomness doesn t help. How could this be? Perhaps we can emulate randomness with pseudo-random generators---deterministic algorithms whose output looks random. What does it mean to look random? A polynomial-time TM can t distinguish them from random. Current research! conp

52 Cryptography! Next time

53 MIT OpenCourseWare J / J Automata, Computability, and Complexity Spring 2011 For information about citing these materials or our Terms of Use, visit:

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