# Lecture 16: Time Complexity and P vs NP

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1 6.045 Lecture 16: Time Complexity and P vs NP 1

2 Time-Bounded Complexity Classes Definition: TIME(t(n)) = { L there is a Turing machine M with time complexity O(t(n)) so that L = L(M) } = { L L is a language decided by a Turing machine with c t(n) + c running time } We showed: A = { 0 k 1 k k 0 } TIME(n log n) Puzzle: Show A TIME( (n log n)/loglog n) 2

3 An Efficient Universal TM Theorem: There is a (one-tape) Turing machine U which takes as input: - the code of an arbitrary TM M - an input string w - and a string of t 1s, t > w such that U(M, w, 1 t ) halts in O( M 2 t 2 ) steps and U accepts (M, w, 1 t ) M accepts w in t steps The Universal TM with a Clock Idea: Make a multi-tape TM U that does the above, and runs in O( M t) steps. Convert to one-tape TM. 3

4 The Time Hierarchy Theorem Intuition: If you get more time to compute, then you can solve strictly more problems. Theorem: For all reasonable f, g : N! N where for all n, g(n) > n 2 f(n) 2, TIME(f(n)) TIME(g(n)) Proof Idea: Diagonalization with a clock Make a TM N that on input M, simulates the TM M on input M for f( M ) steps, then flips the answer. We will show L(N) cannot have time complexity f(n) 4

5 The Time Hierarchy Theorem Theorem: For reasonable f, g where g(n) > n 2 f(n) 2, TIME(f(n)) TIME(g(n)) Proof Sketch: Define a TM N as follows: N(M) = Compute t = f( M ) Run U(M, M, 1 t ) and output the opposite answer. Claim: L(N) does not have time complexity f(n). Proof: Assume N runs in f(n) time, and L(N ) = L(N). By assumption, N (N ) runs in f( N ) time and outputs the opposite answer of U(N, N, 1 f( N ) ) So N (N ) accepts U(N, N, 1 f( N ) ) rejects N (N ) rejects in f( N ) steps [U is universal] This is a contradiction! 5

6 The Time Hierarchy Theorem Theorem: For reasonable f, g where g(n) > n 2 f(n) 2, TIME(f(n)) TIME(g(n)) Proof Sketch: Define a TM N as follows: N(M) = Compute t = f( M ) Run U(M, M, 1 t ) and output the opposite answer. So, L(N) does not have time complexity f(n). For what functions g(n) will N run in O(g(n)) time? 1. Compute t = f( M ) in O(g( M )) time [ reasonable ] 2. Run U(M, M, 1 t ) in O(g( M )) time Recall: U(M, w, 1 t ) halts in O( M 2 t 2 ) steps So set g(n) so that g( M ) > M 2 f( M ) 2 for all n. QED Remark: Time hierarchy also holds for multitape TMs! 6

7 The Time Hierarchy Theorem Theorem: For reasonable f, g where g(n) > n 2 f(n) 2, TIME(f(n)) TIME(g(n)) Proof Sketch: Define a TM N with Σ = 0, 1 as follows: N(0 c 1M) = Compute t = c f( 0 c 1M ) + c Run U(M, 0 c 1M, 1 t ) and output opposite answer. Claim: L(N) does not have time complexity O(f(n)). Proof: Assume N runs in c f(n) + c time, and L(N ) = L(N). Therefore N (0 c 1N ) runs in c f( 0 c 1N ) + c time and outputs the opposite of U(N, 0 c 1N, 1 c f( 0 c1n )+c ) So N (0 c 1N ) accepts U(N, 0 c 1N, 1 c f( 0 c1n )+c ) rejects N (0 c 1N ) rejects in c f( N ) + c time This is a contradiction! 7

8 The Time Hierarchy Theorem Theorem: For reasonable f, g where g(n) > n 2 f(n) 2, TIME(f(n)) TIME(g(n)) Proof Sketch: Define a TM N with Σ = 0, 1 as follows: N(0 c 1M) = Compute t = c f( 0 c 1M ) + c Run U(M, 0 c 1M, 1 t ) and output opposite answer. So, L(N) does not have time complexity f(n). For what functions g(n) will N run in O(g(n)) time? 1. Compute t = c f( 0 c 1M ) + c in O(g( 0 c 1M )) time 2. Simulate U(M, 0 c 1M, 1 t ) in O(g( 0 c 1M )) time Recall: U(M, w, 1 t ) halts in O( M 2 t 2 ) steps Set g(n) s.t. g( 0 c 1M ) > M 2 (c f( 0 c 1M ) + c) 2 for all n. Remark: Time hierarchy also holds for multitape TMs! 8

9 A Better Time Hierarchy Theorem Theorem: For reasonable f, g where g(n) > f(n) log 2 f(n), TIME(f(n)) TIME(g(n)) Corollary: TIME(n) TIME(n 2 ) TIME(n 3 ) There is an infinite hierarchy of increasingly more time-consuming problems Question: Are there important everyday problems that are high up in this time hierarchy? A natural problem that needs exactly n 10 time? THIS IS AN OPEN QUESTION! 9

10 k N P = TIME(n k ) Polynomial Time The analogue of decidability in the world of complexity theory 10

11 The EXTENDED Church-Turing Thesis Everyone s Intuitive Notion of Efficient Algorithms = Polynomial-Time Turing Machines A controversial (dead?) thesis! Counterexamples include n 100 time algorithms, randomized algorithms, quantum algorithms, 11

12 Nondeterminism and NP 12

13 Nondeterministic Turing Machines are just like standard TMs, except: 1. The machine may proceed according to several possible transitions (like an NFA) 2. The machine accepts an input string if there exists an accepting computation history for the machine on the string 13

14 read write move 0 0, R 0 0, R, R q accept, R 0 0, R q reject 14

15 Definition: A nondeterministic TM is a 7-tuple T = (Q, Σ, Γ,, q 0, q accept, q reject ), where: Q is a finite set of states Σ is the input alphabet, where Σ Γ is the tape alphabet, where Γ and Σ Γ : Q Γ 2 (Q Γ {L,R}) q 0 Q is the start state q accept Q is the accept state q reject Q is the reject state, and q reject q accept 15

16 Defining Acceptance for NTMs Let N be a nondeterministic Turing machine An accepting computation history for N on w is a sequence of configurations C 0,C 1,,C t where 1. C 0 is the start configuration q 0 w, 2. C t is an accepting configuration, 3. Each configuration C i yields C i+1 Def. N(w) accepts in t time Such a history exists N has time complexity T(n) if for all n, for all inputs of length n and for all histories, N halts in T(n) time 16

17 Definition: NTIME(t(n)) = { L L is decided by a O(t(n)) time nondeterministic Turing machine } Note: TIME(t(n)) NTIME(t(n)) Is TIME(t(n)) = NTIME(t(n)) for all t(n)? THIS IS AN OPEN QUESTION! What can be done in short NTIME that cannot be done in short TIME? 17

18 Boolean Formulas logical parentheses operations A satisfying assignment is a setting of the variables that makes the formula true = ( x y) z x = 1, y = 1, z = 1 is a satisfying assignment for Boolean variables (0 or 1) (x y) (z x)

19 A Boolean formula is satisfiable if there exists a true/false setting to the variables that makes the formula true YES NO a b c d (x y) x SAT = { is a satisfiable Boolean formula } (How are we encoding formulas? In a reasonable way!) Encoding: takes formula φ of n symbols, and outputs O(n c ) bits Decoding: takes O(n c ) bits and i, and outputs i-th symbol of φ 19

20 A 3cnf-formula has the form: (x 1 x 2 x 3 ) (x 4 x 2 x 5 ) (x 3 x 2 x 1 ) literals clauses Ex: (x 1 x 2 x 1 ) (x 3 x 1 ) (x 3 x 2 x 1 ) (x 1 x 2 x 3 ) ( x 4 x 2 x 1 ) (x 3 x 1 x 1 ) (x 1 x 2 x 3 ) (x 3 x 2 x 1 ) 3SAT = { is a satisfiable 3cnf-formula } 20

21 3SAT = { is a satisfiable 3cnf-formula } Theorem: 3SAT NTIME(n c ) for some constant c > 1 Proof Idea: On input : 1. Check if the formula is in 3cnf 2. For each variable v in, nondeterministically substitute either 0 or 1 in place of v ( x y x ) ( 0 y 0 ) ( 1 y 1 ) ( ) ( ) 3. Evaluate the formula and accept iff is true 21

22 NP = NTIME(n k ) k N Nondeterministic Polynomial Time The analogue of recognizability 22

23 Theorem: L NP There is a constant k and polynomial-time TM V such that L = { x y ϵ Σ* [ y x k and V(x,y) accepts ] } Proof: (1) If L = { x y y x k and V(x,y) accepts } then L NP Given the poly-time TM V, our NP machine for L is: N(x): Nondeterministically guess y. Run V(x,y) (2) If L NP then L = { x y y x k and V(x,y) accepts } Let N be a nondet. poly-time TM that decides L. Define a TM V(x,y) which accepts y encodes an accepting computation history of N on x 23

24 Moral: A language L is in NP if and only if there are polynomial-length proofs for membership in L 3SAT = { y such that is in 3cnf and y is a satisfying assignment to } SAT = { y such that is a Boolean formula and y is a satisfying assignment to } 24

25 NP = Problems with the property that, once you have the answer, it is easy to verify the answer SAT is in NP because a satisfying assignment is a polynomial-length proof that a formula is satisfiable When ϵ SAT, I can prove that fact to you with a short proof you can quickly verify 25

26 The Hamiltonian Path Problem b f e g a d i c h A Hamiltonian path traverses through each node exactly once 26

27 Assume a reasonable encoding of graphs (example: the adjacency matrix is reasonable) HAMPATH = { (G,s,t) G is a directed graph with a Hamiltonian path from s to t } Theorem: HAMPATH NP A Hamiltonian path P in G from s to t is a proof that (G,s,t) is in HAMPATH Given P (as a permutation on the nodes) can easily check that it is a path through all nodes exactly once 27

28 The k-clique Problem a d f c b e g k-clique = complete subgraph on k nodes 28

29 CLIQUE = { (G,k) G is an undirected graph with a k-clique } Theorem: CLIQUE NP A k-clique in G is a proof that (G, k) is in CLIQUE Given a subset S of k nodes from G, we can efficiently check that all possible edges are present between the nodes in S 29

30 A language is in NP if and only if there are polynomial-length proofs for membership in the language P = the problems that can be efficiently solved NP = the problems where proposed solutions can be efficiently verified Is P = NP? can problem solving be automated? 30

31 \$\$\$ P = NP? \$\$\$ 31

32 If P = NP Mathematicians/creators may be out of a job Cryptography as we know it may be impossible no one-way functions! In principle, every aspect of daily life could be efficiently and globally optimized life as we know it would be different Conjecture: P NP 32

33 Next Episode: NP-Complete Problems 33

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