A.Antonopoulos 18/1/2010

Transcription

1 Class DP Basic orems 18/1/2010

2 Class DP Basic orems 1 Class DP 2 Basic orems

3 Class DP Basic orems TSP Versions 1 TSP (D) 2 EXACT TSP 3 TSP COST 4 TSP (1) P (2) P (3) P (4)

4 DP Class Class DP Basic orems A language L is in the class DP if and only if there are two languages L 1 NP and L 2 conp such that L = L 1 L 2. DP is not NP conp! Also, DP is a syntactic class, and so it has complete problems.

5 DP Class Class DP Basic orems A language L is in the class DP if and only if there are two languages L 1 NP and L 2 conp such that L = L 1 L 2. DP is not NP conp! Also, DP is a syntactic class, and so it has complete problems. SAT-UNSAT Given two Boolean expressions φ, φ, both in 3CNF. Is it true that φ is satisfiable and φ is not?

6 Complete for DP Class DP Basic orems

7 Complete for DP orem SAT-UNSAT is DP-complete. Class DP Basic orems

8 Complete for DP Class DP Basic orems orem SAT-UNSAT is DP-complete. Proof Firstly, we have to show it is in DP. So, let: L 1 ={(φ, φ ): φ is satisfiable}. L 2 ={(φ, φ ): φ is unsatisfiable}. It is easy to see, L 1 NP and L 2 conp, thus L L 1 L 2 DP. For completeness, let L DP. We have to show that L P SAT-UNSAT. L DP L = L 1 L 2, L 1 NP and L 2 conp. SAT NP-complete R 1 :L 1 P SAT and R 2 :L 2 P SAT. Hence, L P SAT-UNSAT,by R(x) = (R 1 (x), R 2 (x))

9 Complete for DP Class DP Basic orems orem EXACT TSP is DP-complete. Proof EXACT TSP DP, by L 1 TSP NP and L 2 TSP COMPLEMENT conp Completeness: we ll show that SAT-UNSAT P EXACT TSP. 3SAT P HP: (φ, φ ) (G, G ) Broken Hamilton Path (2 node-disjoint paths that cover all nodes) Almost Satisfying Truth Assignement (satisfies all clauses except for one)

10 Complete for DP Class DP Basic orems Proof We define distances: 1 If (i, j) E(G) or E(G ): d(i, j) 1 2 If (i, j) / E(G), but i and j V(G): d(i, j) 2 3 Otherwise: d(i, j) 4 Let n be the size of the graph. 1 If φ and φ satisfiable, then optcost = n 2 If φ and φ unsatisfiable, then optcost = n If φ satisfiable and φ not, then optcost = n If φ satisfiable and φ not, then optcost = n + 1 yes instance of SAT-UNSAT optcost = n + 2 Let B n + 2!

11 Other DP-complete problems Class DP Basic orems Also: CRITICAL SAT: Given a Boolean expression φ, is it true that it s unsatisfiable, but deleting any clause makes it satisfiable? CRITICAL HAMILTON PATH: Given a graph, is it true that it has no Hamilton path, but addition of any edge creates a Hamilton path? CRITICAL 3-COLORABILITY: Given a graph, is it true that it is not 3-colorable, but deletion of any node makes it 3-colorable? are DP-complete!

12 Classes P NP and FP NP Class DP Basic orems Alternative DP DP is the class of languages that can be decided by an oracle machine which makes 2 queries to a SAT oracle, and accepts iff the 1st answer is yes, and the 2nd is no. P SAT is the class of languages decided in pol time with a SAT oracle. number of queries Queries computed adaptively SAT NP-complete P SAT =P NP FP NP is the class of functions that can be computed by a pol-time TM with a SAT oracle. Goal: MAX OUTPUT P MAX-WEIGHT SAT P SAT

13 FP NP -complete Class DP Basic orems MAX OUTPUT Given NTM N, with input 1 n, which halts after O(n),with output a string of length n. Which is the largest output,of any computation of N on 1 n?

14 FP NP -complete Class DP Basic orems MAX OUTPUT Given NTM N, with input 1 n, which halts after O(n),with output a string of length n. Which is the largest output,of any computation of N on 1 n?

15 FP NP -complete Class DP Basic orems MAX OUTPUT Given NTM N, with input 1 n, which halts after O(n),with output a string of length n. Which is the largest output,of any computation of N on 1 n? orem MAX OUTPUT is FP NP -complete. Proof MAX OUTPUT FP NP. Let F : Σ Σ FP NP pol-time TM M?, s.t. M SAT (x) = F (x) We ll show: F MAX OUTPUT! Reductions R and S (log space computable) s.t.: x, R(x) is a instance of MAX OUTPUT S(max output of R(x)) F (x)

16 FP NP -complete Class DP Basic orems Proof NTM N: Let n = p 2 ( x ), p( ), is the pol bound of SAT. N(1 n ) generates x on a string. M SAT query state (φ 1 ): If z 1 = 0 (φ 1 unsat), then continue from q NO. If z 1 = 1 (φ 1 sat), then guess assignment T 1 : If test succeeds, continue from q YES. If test fails, output=0 n and halt. (Unsuccessful computation) Continue to all guesses (z i ), and halt, with output= z 1 z }{{} n (Successful computation)

17 FP NP -complete Class DP Basic orems Proof We claim that the successful computation that outputs the largest integer, correspond to a correct simulation: Let j the smallest integer,s.t.: z j = 0, while φ j was satisfiable. n, another successful computation of N, s.t.: z j = 1. computations agree to the first j 1 digits, the 2 nd represents a larger number. S part: F (x) can be read off the end of the largest output of N.

18 FP NP -complete Class DP Basic orems MAX-WEIGHT SAT Given a set of clauses, each with an integer weight, find the truth assignment that satisfies a set of clauses with the most total weight.

19 FP NP -complete Class DP Basic orems MAX-WEIGHT SAT Given a set of clauses, each with an integer weight, find the truth assignment that satisfies a set of clauses with the most total weight. orem MAX-WEIGHT SAT is FP NP -complete. Proof MAX-WEIGHT SAT is in FP NP : By binary search, and a SAT oracle, we can find the largest possible total weight of satisfied clauses, and then, by setting the variables 1-1, the truth assignment that achieves it. MAX OUTPUT MAX-WEIGHT SAT:

20 FP NP -complete Class DP Basic orems Proof NTMN(1 n ) φ(n, m): Any satisfying truth assignment of φ(n, m) legal comp. of N(1 n ) Clauses are given a huge weight (2 n ), so that any t.a. that aspires to be optimum satisfy all clauses of φ(n, m). Add more clauses: (y i ): i = 1,..n with weight 2 n i. Now, optimum t.a. must not represent any legal computation, but this which produces the largest possible output value. S part: From optimum t.a. of the resulting expression (or the weight), we can recover the optimum output of N(1 n ).

21 FP NP -complete Class DP Basic orems And the main result: orem TSP is FP NP -complete.

22 FP NP -complete Class DP Basic orems And the main result: orem TSP is FP NP -complete. Corollary TSP COST is FP NP -complete.

23 FP NP -complete Class DP Basic orems Figure: overall construction (17-2)

24 NP[log n] Class P Class DP Basic orems P NP[logn] is the class of all languages decided by a polynomial time oracle machine, which on input x asks a total of O(log x ) SAT queries. FP NP[logn] is the corresponding class of functions.

25 NP[log n] Class P Class DP Basic orems P NP[logn] is the class of all languages decided by a polynomial time oracle machine, which on input x asks a total of O(log x ) SAT queries. FP NP[logn] is the corresponding class of functions. CLIQUE SIZE Given a graph, determine the size of his largest clique.

26 NP[log n] Class P Class DP Basic orems P NP[logn] is the class of all languages decided by a polynomial time oracle machine, which on input x asks a total of O(log x ) SAT queries. FP NP[logn] is the corresponding class of functions. CLIQUE SIZE Given a graph, determine the size of his largest clique. orem CLIQUE SIZE is FP NP[logn] -complete.

27 Conclusion Class DP Basic orems 1 TSP (D) is NP-complete. 2 EXACT TSP is DP-complete. 3 TSP COST is FP NP -complete. 4 TSP is FP NP -complete. And now, P NP NP NP? Oracles for NP NP?

28 Class DP Basic orems 1 Class DP 2 Basic orems

29 Class DP Basic orems 0 P = Σ 0 P = Π 0 P = P i+1 P = P Σ i P Σ i+1 P = NP Σ i P Π i+1 P = conp Σ i P PH i 0 Σ i P

30 Class DP Basic orems 0 P = Σ 0 P = Π 0 P = P i+1 P = P Σ i P Σ i+1 P = NP Σ i P Π i+1 P = conp Σ i P PH i 0 Σ i P

31 Class DP Basic orems 0 P = Σ 0 P = Π 0 P = P i+1 P = P Σ i P Σ i+1 P = NP Σ i P Π i+1 P = conp Σ i P PH i 0 Σ i P Σ 0 P = P 1 P = P, Σ 1 P = NP, Π 1 P = conp 2 P = P NP, Σ 2 P = NP NP, Π 2 P = conp NP

32 Basic orems Class DP Basic orems orem Let L be a language, and i 1. L Σ i P iff there is a polynomially balanced relation R such that the language {x; y : (x, y) R} is in Π i 1 P and Proof (by Induction) L = {x : y, s.t. : (x, y) R} For i = 1 {x; y : (x, y) R} P,so L = {x y : (x, y) R} NP For i > 1 If R Π i 1 P, we must show that L Σ i P NTM with Σ i 1 P oracle: NTM(x) guesses a y and asks Σ i 1 P oracle whether (x, y) / R.

33 Basic orems Class DP Basic orems Proof If L Σ i P, we must show the existence or R. L Σ i P NTM M K, K Σ i 1 P, which decides L. K Σ i 1 P S Π i 2 P : (z K w : (z, w) S) We must describe a relation R (we know: x L accepting comp of M K (x)) Query Steps: yes z i has a cerfificate w i st (z i, w i ) S. So, R(x) = (x, y) R iff yrecords an accepting computation of M? on x, together with a certificate w i for each yes query z i in the computation. We must show {x; y : (x, y) R} Π i 1 P.

34 Basic orems Class DP Basic orems Corollary Let L be a language, and i 1. L Π i P iff there is a polynomially balanced relation R such that the language {x; y : (x, y) R} is in Σ i 1 P and Corollary L = {x : y, y x k, s.t. : (x, y) R} Let L be a language, and i 1. L Σ i P iff there is a polynomially balanced, polynomially-time decicable (i + 1)-ary relation R such that: L = {x : y 1 y 2 y 3...Qy i, s.t. : (x, y 1,..., y i ) R} where the i th quantifier Q is, if i is even, and, if i is odd.

35 Basic orems Class DP Basic orems orem If for some i 1, Σ i P = Π i P, then for all j > i: Σ j P = Π j P = j P = Σ i P Or, the polynomial hierarchy collapses to the i th level. Proof It suffices to show that: Σ i P = Π i P Σ i+1 P = Σ i P Let L Σ i+1 P R Π i P: L = {x y : (x, y) R} Since Π i P = Σ i P R Σ i P (x, y) R z : (x, y, z) S, S Π i 1 P. Thus, x L y; z : (x, y, z) S, S Π i 1 P, which means L Σ i P.

36 Basic orems Class DP Basic orems Corollary If P=NP, or even NP=coNP, the collapses to the first level.

37 Basic orems Class DP Basic orems Corollary If P=NP, or even NP=coNP, the collapses to the first level. MINIMUM CIRCUIT Given a Boolean Circuit C, is it true that there is no circuit with fewer gates that computes the same Boolean function

38 Basic orems Class DP Basic orems Corollary If P=NP, or even NP=coNP, the collapses to the first level. MINIMUM CIRCUIT Given a Boolean Circuit C, is it true that there is no circuit with fewer gates that computes the same Boolean function

39 Basic orems Class DP Basic orems Corollary If P=NP, or even NP=coNP, the collapses to the first level. MINIMUM CIRCUIT Given a Boolean Circuit C, is it true that there is no circuit with fewer gates that computes the same Boolean function MINIMUM CIRCUIT is in Π 2 P, and not known to be in any class below that. It is open whether MINIMUM CIRCUIT is Π 2 P-complete.

40 Basic orems Class DP Basic orems QSAT i Given expression φ, with Boolean variables partitioned into i sets X i,is φ satisfied by the overall truth assignment of the expression: X 1 X 2 X 3...QX i φ, where Q is if i is odd, and if i is even.

41 Basic orems Class DP Basic orems QSAT i Given expression φ, with Boolean variables partitioned into i sets X i,is φ satisfied by the overall truth assignment of the expression: X 1 X 2 X 3...QX i φ, where Q is if i is odd, and if i is even. orem For all i 1 QSAT i is Σ i P-complete.

42 Basic orems Class DP Basic orems orem If there is a PH-complete problem, then the polynomial hierarchy collapses to some finite level. Proof Let L is PH-complete. Since L PH, i 0 : L Σ i P. But any L Σ i+1 P reduces to L. Since PH is closed under reductions, we imply that L Σ i P, so Σ i P = Σ i+1 P.

43 Basic orems Class DP Basic orems orem If there is a PH-complete problem, then the polynomial hierarchy collapses to some finite level. Proof Let L is PH-complete. Since L PH, i 0 : L Σ i P. But any L Σ i+1 P reduces to L. Since PH is closed under reductions, we imply that L Σ i P, so Σ i P = Σ i+1 P. orem PH PSPACE

44 Basic orems Class DP Basic orems orem If there is a PH-complete problem, then the polynomial hierarchy collapses to some finite level. Proof Let L is PH-complete. Since L PH, i 0 : L Σ i P. But any L Σ i+1 P reduces to L. Since PH is closed under reductions, we imply that L Σ i P, so Σ i P = Σ i+1 P. orem PH PSPACE

45 Basic orems Class DP Basic orems orem If there is a PH-complete problem, then the polynomial hierarchy collapses to some finite level. Proof Let L is PH-complete. Since L PH, i 0 : L Σ i P. But any L Σ i+1 P reduces to L. Since PH is closed under reductions, we imply that L Σ i P, so Σ i P = Σ i+1 P. orem PH PSPACE PH? = PSPACE (Open). If it was, then PH has complete problems, so it collapses to some finite level.

46 Class DP Basic orems orem BPP Σ 2 P Π 2 P Proof Because cobpp = BPP,we prove only BPP Σ 2 P. Let L BPP (L is accepted by clear majority ). For x = n, let A(x) {0, 1} p(n) be the set of accepting computations. We have: x L A(x) 2 p(n) ( n ) x / L A(x) 2 p(n) ( 1 2 n ) Let U be the set of all bit strings of length p(n). For a, b U, let a b be the XOR: a b = c c b = a, so b is 1-1.

47 Class DP Basic orems Proof For t U, A(x) t = {a t : a A(x)} (translation of A(x) by t). We imply that: A(x) t = A(x) If x L, consider a random (drawing p 2 (n) bits) sequence of translations: t 1, t 2,.., t p(n) U. For b U, these translations cover b, if b A(x) t j, j p(n). b A(x) t j b t j A(x) Pr[b / A(x) t j ]= 1 2 n Pr[b is not covered by any t j ]=2 np(n) Pr[ point that is not covered] 2 np(n) U = 2 (n 1)p(n)

48 Class DP Basic orems Proof So,T = (t 1,.., t p(n) ) has a positive probability that it covers all of U. If x / L, A(x) is exp small,and (for large n) there s not T that cover all U. (x L) ( T that cover all U) So, L = {x (T {0, 1} p2 (n) ) (b U) (j p(n)) : b t j A(x)} which is precisely the form of languages in Σ 2 P. last existential quantifier ( (j p(n))...) affects only polynomially many possibilities,so it doesn t count (can by tested in polynomial time by trying all t j s).