True quantified Boolean formula




In computational complexity theory, the language TQBF is a formal language consisting of the true quantified Boolean formulas. A (fully) quantified Boolean formula is a formula in quantified propositional logic where every variable is quantified (or bound), using either existential or universal quantifiers, at the beginning of the sentence. Such a formula is equivalent to either true or false (since there are no free variables). If such a formula evaluates to true, then that formula is in the language TQBF. It is also known as QSAT (Quantified SAT).




Contents






  • 1 Overview


  • 2 Prenex normal form


  • 3 Solving


  • 4 PSPACE-completeness


  • 5 Miscellany


  • 6 Notes and references


  • 7 See also


  • 8 External links





Overview


In computational complexity theory, the quantified Boolean formula problem (QBF) is a generalization of the Boolean satisfiability problem in which both existential quantifiers and universal quantifiers can be applied to each variable. Put another way, it asks whether a quantified sentential form over a set of Boolean variables is true or false. For example, the following is an instance of QBF:


∀x ∃y ∃z ((x∨z)∧y){displaystyle forall x exists y exists z ((xlor z)land y)}forall x exists y exists z ((x lor z) land y)

QBF is the canonical complete problem for PSPACE, the class of problems solvable by a deterministic or nondeterministic Turing machine in polynomial space and unlimited time.[1] Given the formula in the form of an abstract syntax tree, the problem can be solved easily by a set of mutually recursive procedures which evaluate the formula. Such an algorithm uses space proportional to the height of the tree, which is linear in the worst case, but uses time exponential in the number of quantifiers.


Provided that MA ⊊ PSPACE, which is widely believed, QBF cannot be solved, nor can a given solution even be verified, in either deterministic or probabilistic polynomial time (in fact, unlike the satisfiability problem, there's no known way to specify a solution succinctly). It can be solved using an alternating Turing machine in linear time, since AP = PSPACE, where AP is the class of problems alternating machines can solve in polynomial time.[2]


When the seminal result IP = PSPACE was shown (see interactive proof system), it was done by exhibiting an interactive proof system that could solve QBF by solving a particular arithmetization of the problem.[3]


QBF formulas have a number of useful canonical forms. For example, it can be shown that there is a polynomial-time many-one reduction that will move all quantifiers to the front of the formula and make them alternate between universal and existential quantifiers. There is another reduction that proved useful in the IP = PSPACE proof where no more than one universal quantifier is placed between each variable's use and the quantifier binding that variable. This was critical in limiting the number of products in certain subexpressions of the arithmetization.



Prenex normal form


A fully quantified Boolean formula can be assumed to have a very specific form, called prenex normal form. It has two basic parts: a portion containing only quantifiers and a portion containing an unquantified Boolean formula usually denoted as ϕ{displaystyle displaystyle phi }displaystyle phi. If there are n{displaystyle displaystyle n}displaystyle n Boolean variables, the entire formula can be written as


∃x1∀x2∃x3⋯Qnxnϕ(x1,x2,x3,…,xn){displaystyle displaystyle exists x_{1}forall x_{2}exists x_{3}cdots Q_{n}x_{n}phi (x_{1},x_{2},x_{3},dots ,x_{n})}displaystyle exists x_1 forall x_2 exists x_3 cdots Q_n x_n phi(x_1, x_2, x_3, dots, x_n)

where every variable falls within the scope of some quantifier. By introducing dummy variables, any formula in prenex normal form can be converted into a sentence where existential and universal quantifiers alternate. Using the dummy variable y1{displaystyle displaystyle y_{1}}displaystyle y_1,


∃x1∃x2ϕ(x1,x2)↦∃x1∀y1∃x2ϕ(x1,x2){displaystyle displaystyle exists x_{1}exists x_{2}phi (x_{1},x_{2})quad mapsto quad exists x_{1}forall y_{1}exists x_{2}phi (x_{1},x_{2})}displaystyle exists x_1 exists x_2 phi(x_1, x_2) quad mapsto quad<br /> exists x_1 forall y_1 exists x_2 phi(x_1, x_2)

The second sentence has the same truth value but follows the restricted syntax. Assuming fully quantified Boolean formulas to be in prenex normal form is a frequent feature of proofs.



Solving


There is a simple recursive algorithm for determining whether a QBF is in TQBF (i.e. is true). Given some QBF


Q1x1Q2x2⋯Qnxnϕ(x1,x2,…,xn).{displaystyle Q_{1}x_{1}Q_{2}x_{2}cdots Q_{n}x_{n}phi (x_{1},x_{2},dots ,x_{n}).}Q_1 x_1 Q_2 x_2 cdots Q_n x_n phi(x_1, x_2, dots, x_n).

If the formula contains no quantifiers, we can just return the formula. Otherwise, we take off the first quantifier and check both possible values for the first variable:



A=Q2x2⋯Qnxnϕ(0,x2,…,xn),{displaystyle A=Q_{2}x_{2}cdots Q_{n}x_{n}phi (0,x_{2},dots ,x_{n}),}A = Q_2 x_2 cdots Q_n x_n phi(0, x_2, dots, x_n),

B=Q2x2⋯Qnxnϕ(1,x2,…,xn).{displaystyle B=Q_{2}x_{2}cdots Q_{n}x_{n}phi (1,x_{2},dots ,x_{n}).}B = Q_2 x_2 cdots Q_n x_n phi(1, x_2, dots, x_n).


If Q1=∃{displaystyle Q_{1}=exists }Q_1 = exists, then return A∨B{displaystyle Alor B}Alor B. If Q1=∀{displaystyle Q_{1}=forall }Q_1 = forall, then return A∧B{displaystyle Aland B}Aland B.


How fast does this algorithm run?
For every quantifier in the initial QBF, the algorithm makes two recursive calls on only a linearly smaller subproblem. This gives the algorithm an exponential runtime O(2n).


How much space does this algorithm use?
Within each invocation of the algorithm, it needs to store the intermediate results of computing A and B. Every recursive call takes off one quantifier, so the total recursive depth is linear in the number of quantifiers. Formulas that lack quantifiers can be evaluated in space logarithmic in the number of variables. The initial QBF was fully quantified, so there are at least as many quantifiers as variables. Thus, this algorithm uses O(n + log n) = O(n) space. This makes the TQBF language part of the PSPACE complexity class.



PSPACE-completeness


The TQBF language serves in complexity theory as the canonical PSPACE-complete problem. Being PSPACE-complete means that a language is in PSPACE and that the language is also PSPACE-hard. The algorithm above shows that TQBF is in PSPACE.
Showing that TQBF is PSPACE-hard requires showing that any language in the complexity class PSPACE can be reduced to TQBF in polynomial time. I.e.,


∀L∈PSPACE,L≤pTQBF.{displaystyle forall Lin {mathsf {PSPACE}},Lleq _{p}mathrm {TQBF} .}{displaystyle forall Lin {mathsf {PSPACE}},Lleq _{p}mathrm {TQBF} .}

This means that, for a PSPACE language L, whether an input x is in L can be decided by checking whether f(x){displaystyle f(x)}f(x) is in TQBF, for some function f that is required to run in polynomial time (relative to the length of the input) Symbolically,


x∈L⟺f(x)∈TQBF.{displaystyle xin Liff f(x)in mathrm {TQBF} .}{displaystyle xin Liff f(x)in mathrm {TQBF} .}

Proving that TQBF is PSPACE-hard, requires specification of f.


So, suppose that L is a PSPACE language. This means that L can be decided by a polynomial space deterministic Turing machine (DTM). This is very important for the reduction of L to TQBF, because the configurations of any such Turing Machine can be represented as Boolean formulas, with Boolean variables representing the state of the machine as well as the contents of each cell on the Turing Machine tape, with the position of the Turing Machine head encoded in the formula by the formula's ordering. In particular, our reduction will use the variables c1{displaystyle c_{1}}c_{1} and c2{displaystyle c_{2}}c_{2}, which represent two possible configurations of the DTM for L, and a natural number t, in constructing a QBF ϕc1,c2,t{displaystyle phi _{c_{1},c_{2},t}}phi_{c_1,c_2,t} which is true if and only if the DTM for L can go from the configuration encoded in c1{displaystyle c_{1}}c_{1} to the configuration encoded in c2{displaystyle c_{2}}c_{2} in no more than t steps. The function f, then, will construct from the DTM for L a QBF ϕcstart,caccept,T{displaystyle phi _{c_{start},c_{accept},T}}phi_{c_{start},c_{accept},T}, where cstart{displaystyle c_{start}}c_{start} is the DTM's starting configuration, caccept{displaystyle c_{accept}}c_{accept} is the DTM's accepting configuration, and T is the maximum number of steps the DTM could need to move from one configuration to the other. We know that T = O(exp(n)), where n is the length of the input, because this bounds the total number of possible configurations of the relevant DTM. Of course, it cannot take the DTM more steps than there are possible configurations to reach caccept{displaystyle c_{mathrm {accept} }}c_mathrm{accept} unless it enters a loop, in which case it will never reach caccept{displaystyle c_{mathrm {accept} }}c_mathrm{accept} anyway.


At this stage of the proof, we have already reduced the question of whether an input formula w (encoded, of course, in cstart{displaystyle c_{start}}c_{start}) is in L to the question of whether the QBF ϕcstart,caccept,T{displaystyle phi _{c_{start},c_{accept},T}}phi_{c_{start},c_{accept},T}, i.e., f(w){displaystyle f(w)}f(w), is in TQBF. The remainder of this proof proves that f can be computed in polynomial time.


For t=1{displaystyle t=1}t=1, computation of ϕc1,c2,t{displaystyle phi _{c_{1},c_{2},t}}phi_{c_1,c_2,t} is straightforward—either one of the configurations changes to the other in one step or it does not. Since the Turing Machine that our formula represents is deterministic, this presents no problem.


For t>1{displaystyle t>1}t>1, computation of ϕc1,c2,t{displaystyle phi _{c_{1},c_{2},t}}phi_{c_1,c_2,t} involves a recursive evaluation, looking for a so-called "middle point" m1{displaystyle m_{1}}m_{1}. In this case, we rewrite the formula as follows:


ϕc1,c2,t=∃m1(ϕc1,m1,⌈t/2⌉∧ϕm1,c2,⌈t/2⌉).{displaystyle phi _{c_{1},c_{2},t}=exists m_{1}(phi _{c_{1},m_{1},lceil t/2rceil }wedge phi _{m_{1},c_{2},lceil t/2rceil }).}phi_{c_1,c_2,t}=exists m_1(phi_{c_1,m_1,lceil t/2rceil}wedgephi_{m_1,c_2,lceil t/2rceil}).

This converts the question of whether c1{displaystyle c_{1}}c_{1} can reach c2{displaystyle c_{2}}c_{2} in t steps to the question of whether c1{displaystyle c_{1}}c_{1} reaches a middle point m1{displaystyle m_{1}}m_{1} in t/2{displaystyle t/2}t/2 steps, which itself reaches c2{displaystyle c_{2}}c_{2} in t/2{displaystyle t/2}t/2 steps. The answer to the latter question of course gives the answer to the former.


Now, t is only bounded by T, which is exponential (and so not polynomial) in the length of the input. Additionally, each recursive layer virtually doubles the length of the formula. (The variable m1{displaystyle m_{1}}m_{1} is only one midpoint—for greater t, there are more stops along the way, so to speak.) So the time required to recursively evaluate ϕc1,c2,t{displaystyle phi _{c_{1},c_{2},t}}phi_{c_1,c_2,t} in this manner could be exponential as well, simply because the formula could become exponentially large. This problem is solved by universally quantifying using variables c3{displaystyle c_{3}}c_{3} and c4{displaystyle c_{4}}c_4 over the configuration pairs (e.g., {(c1,m1),(m1,c2)}{displaystyle {(c_{1},m_{1}),(m_{1},c_{2})}}{ (c_1,m_1),(m_1,c_2)}), which prevents the length of the formula from expanding due to recursive layers. This yields the following interpretation of ϕc1,c2,t{displaystyle phi _{c_{1},c_{2},t}}phi_{c_1,c_2,t}:


ϕc1,c2,t=∃m1∀(c3,c4)∈{(c1,m1),(m1,c2)}(ϕc3,c4,⌈t/2⌉).{displaystyle phi _{c_{1},c_{2},t}=exists m_{1}forall (c_{3},c_{4})in {(c_{1},m_{1}),(m_{1},c_{2})}(phi _{c_{3},c_{4},lceil t/2rceil }).}phi_{c_1,c_2,t}=exists m_1forall (c_3,c_4)in { (c_1,m_1),(m_1,c_2)}(phi_{c_3,c_4,lceil t/2rceil}).

This version of the formula can indeed be computed in polynomial time, since any one instance of it can be computed in polynomial time. The universally quantified ordered pair simply tells us that whichever choice of (c3,c4){displaystyle (c_{3},c_{4})}(c_3,c_4) is made, ϕc1,c2,t⟺ϕc3,c4,⌈t/2⌉{displaystyle phi _{c_{1},c_{2},t}iff phi _{c_{3},c_{4},lceil t/2rceil }}phi_{c_1,c_2,t}iffphi_{c_3,c_4,lceil t/2rceil}.


Thus, ∀L∈PSPACE,L≤pTQBF{displaystyle forall Lin {mathsf {PSPACE}},Lleq _{p}mathrm {TQBF} }{displaystyle forall Lin {mathsf {PSPACE}},Lleq _{p}mathrm {TQBF} }, so TQBF is PSPACE-hard. Together with the above result that TQBF is in PSPACE, this completes the proof that TQBF is a PSPACE-complete language.


(This proof follows Sipser 2006 pp. 310–313 in all essentials. Papadimitriou 1994 also includes a proof.)



Miscellany


  • One important subproblem in TQBF is the Boolean satisfiability problem. In this problem, you wish to know whether a given Boolean formula ϕ{displaystyle phi }phi can be made true with some assignment of variables. This is equivalent to the TQBF using only existential quantifiers:

∃x1⋯∃xnϕ(x1,…,xn){displaystyle exists x_{1}cdots exists x_{n}phi (x_{1},ldots ,x_{n})}exists x_1 cdots exists x_n phi(x_1, ldots, x_n)

This is also an example of the larger result NP ⊆{displaystyle subseteq }subseteq PSPACE which follows directly from the observation that a polynomial time verifier for a proof of a language accepted by a NTM (Non-deterministic Turing machine) requires polynomial space to store the proof.

  • Any class in the polynomial hierarchy (PH) has TQBF as a hard problem. In other words, for the class comprising all languages L for which there exists a poly-time TM V, a verifier, such that for all input x and some constant i,

x∈L⇔∃y1∀y2⋯Qiyi V(x,y1,y2,…,yi) = 1{displaystyle xin LLeftrightarrow exists y_{1}forall y_{2}cdots Q_{i}y_{i} V(x,y_{1},y_{2},dots ,y_{i}) = 1}x in L Leftrightarrow exists y_1 forall y_2 cdots Q_i y_i V(x,y_1,y_2,dots,y_i) = 1

which has a specific QBF formulation that is given as


∃ϕ{displaystyle exists phi }exists phi such that ∃x1→∀x2→⋯Qixi→ ϕ(x1→,x2→,…,xi→) = 1{displaystyle exists {vec {x_{1}}}forall {vec {x_{2}}}cdots Q_{i}{vec {x_{i}}} phi ({vec {x_{1}}},{vec {x_{2}}},dots ,{vec {x_{i}}}) = 1}exists vec{x_1} forall vec{x_2} cdots Q_i vec{x_i} phi(vec{x_1},vec{x_2},dots,vec{x_i}) = 1

where the xi→{displaystyle {vec {x_{i}}}}vec{x_i}'s are vectors of Boolean variables.


  • It is important to note that while TQBF the language is defined as the collection of true quantified Boolean formulas, the abbreviation TQBF is often used (even in this article) to stand for a totally quantified Boolean formula, often simply called a QBF (quantified Boolean formula, understood as "fully" or "totally" quantified). It is important to distinguish contextually between the two uses of the abbreviation TQBF in reading the literature.

  • A TQBF can be thought of as a game played between two players, with alternating moves. Existentially quantified variables are equivalent to the notion that a move is available to a player at a turn. Universally quantified variables mean that the outcome of the game does not depend on what move a player makes at that turn. Also, a TQBF whose first quantifier is existential corresponds to a formula game in which the first player has a winning strategy.

  • A TQBF for which the quantified formula is in 2-CNF may be solved in linear time, by an algorithm involving strong connectivity analysis of its implication graph. The 2-satisfiability problem is a special case of TQBF for these formulas, in which every quantifier is existential.[4][5]

  • There is a systematic treatment of restricted versions of quantified boolean formulas (giving Schaefer-type classifications) provided in an expository paper by Hubie Chen.[6]



Notes and references




  1. ^ M. Garey & D. Johnson (1979). Computers and Intractability: A Guide to the Theory of NP-Completeness. W. H. Freeman, San Francisco, California. ISBN 0-7167-1045-5..mw-parser-output cite.citation{font-style:inherit}.mw-parser-output .citation q{quotes:"""""""'""'"}.mw-parser-output .citation .cs1-lock-free a{background:url("//upload.wikimedia.org/wikipedia/commons/thumb/6/65/Lock-green.svg/9px-Lock-green.svg.png")no-repeat;background-position:right .1em center}.mw-parser-output .citation .cs1-lock-limited a,.mw-parser-output .citation .cs1-lock-registration a{background:url("//upload.wikimedia.org/wikipedia/commons/thumb/d/d6/Lock-gray-alt-2.svg/9px-Lock-gray-alt-2.svg.png")no-repeat;background-position:right .1em center}.mw-parser-output .citation .cs1-lock-subscription a{background:url("//upload.wikimedia.org/wikipedia/commons/thumb/a/aa/Lock-red-alt-2.svg/9px-Lock-red-alt-2.svg.png")no-repeat;background-position:right .1em center}.mw-parser-output .cs1-subscription,.mw-parser-output .cs1-registration{color:#555}.mw-parser-output .cs1-subscription span,.mw-parser-output .cs1-registration span{border-bottom:1px dotted;cursor:help}.mw-parser-output .cs1-ws-icon a{background:url("//upload.wikimedia.org/wikipedia/commons/thumb/4/4c/Wikisource-logo.svg/12px-Wikisource-logo.svg.png")no-repeat;background-position:right .1em center}.mw-parser-output code.cs1-code{color:inherit;background:inherit;border:inherit;padding:inherit}.mw-parser-output .cs1-hidden-error{display:none;font-size:100%}.mw-parser-output .cs1-visible-error{font-size:100%}.mw-parser-output .cs1-maint{display:none;color:#33aa33;margin-left:0.3em}.mw-parser-output .cs1-subscription,.mw-parser-output .cs1-registration,.mw-parser-output .cs1-format{font-size:95%}.mw-parser-output .cs1-kern-left,.mw-parser-output .cs1-kern-wl-left{padding-left:0.2em}.mw-parser-output .cs1-kern-right,.mw-parser-output .cs1-kern-wl-right{padding-right:0.2em}


  2. ^ A. Chandra, D. Kozen, and L. Stockmeyer (1981). "Alternation". Journal of the ACM. 28 (1): 114–133. doi:10.1145/322234.322243.CS1 maint: Multiple names: authors list (link)


  3. ^ Adi Shamir (1992). "Ip = Pspace". Journal of the ACM. 39 (4): 869–877. doi:10.1145/146585.146609.


  4. ^ Krom, Melven R. (1967). "The Decision Problem for a Class of First-Order Formulas in Which all Disjunctions are Binary". Zeitschrift für Mathematische Logik und Grundlagen der Mathematik. 13: 15–20. doi:10.1002/malq.19670130104.


  5. ^ Aspvall, Bengt; Plass, Michael F.; Tarjan, Robert E. (1979). "A linear-time algorithm for testing the truth of certain quantified boolean formulas" (PDF). Information Processing Letters. 8 (3): 121–123. doi:10.1016/0020-0190(79)90002-4.


  6. ^ Chen, Hubie (December 2009). "A Rendezvous of Logic, Complexity, and Algebra". ACM Computing Surveys. ACM. 42 (1): 1. arXiv:cs/0611018. doi:10.1145/1592451.1592453.



  • Fortnow & Homer (2003) provides some historical background for PSPACE and TQBF.

  • Zhang (2003) provides some historical background of Boolean formulas.

  • Arora, Sanjeev. (2001). COS 522: Computational Complexity. Lecture Notes, Princeton University. Retrieved October 10, 2005.

  • Fortnow, Lance & Steve Homer. (2003, June). A short history of computational complexity. The Computational Complexity Column, 80. Retrieved October 9, 2005.

  • Papadimitriou, C. H. (1994). Computational Complexity. Reading: Addison-Wesley.

  • Sipser, Michael. (2006). Introduction to the Theory of Computation. Boston: Thomson Course Technology.

  • Zhang, Lintao. (2003). Searching for truth: Techniques for satisfiability of boolean formulas. Retrieved October 10, 2005.



See also




  • Cook–Levin theorem, stating that SAT is NP-complete

  • Generalized geography



External links



  • The Quantified Boolean Formulas Library (QBFLIB)

  • International Workshop on Quantified Boolean Formulas


  • DepQBF - a search-based solver for quantified Boolean formulas


  • CAQE - a CEGAR-based solver for quantified Boolean formulas; winner of the recent editions of the yearly competition of QBF solvers QBFEVAL.


  • CADET - a solver for quantified Boolean formulas restricted to one quantifier alternation with the ability to compute Skolem functions




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