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Limits of Preprocessing

by   Stefan Szeider, et al.

We present a first theoretical analysis of the power of polynomial-time preprocessing for important combinatorial problems from various areas in AI. We consider problems from Constraint Satisfaction, Global Constraints, Satisfiability, Nonmonotonic and Bayesian Reasoning. We show that, subject to a complexity theoretic assumption, none of the considered problems can be reduced by polynomial-time preprocessing to a problem kernel whose size is polynomial in a structural problem parameter of the input, such as induced width or backdoor size. Our results provide a firm theoretical boundary for the performance of polynomial-time preprocessing algorithms for the considered problems.


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1 Introduction

Many important computational problems that arise in various areas of AI are intractable. Nevertheless, AI research was very successful in developing and implementing heuristic solvers that work well on real-world instances. An important component of virtually every solver is a powerful polynomial-time preprocessing procedure that reduces the problem input. For instance, preprocessing techniques for the propositional satisfiability problem are based on Boolean Constraint Propagation (see, e.g., EenBiere05, EenBiere05), CSP solvers make use of various local consistency algorithms that filter the domains of variables (see, e.g., Bessiere06, Bessiere06); similar preprocessing methods are used by solvers for Nonmonotonic and Bayesian reasoning problems (see, e.g., GebserKaufmannNeumannSchaub08, GebserKaufmannNeumannSchaub08, BoltGaag06, BoltGaag06, respectively).

Until recently, no provable performance guarantees for polynomial-time preprocessing methods have been obtained, and so preprocessing was only subject of empirical studies. A possible reason for the lack of theoretical results is a certain inadequacy of the P vs NP framework for such an analysis: if we could reduce in polynomial time an instance of an NP-hard problem just by one bit, then we can solve the entire problem in polynomial time by repeating the reduction step a polynomial number of times, and follows.

With the advent of parameterized complexity [Downey, Fellows, and Stege1999], a new theoretical framework became available that provides suitable tools to analyze the power of preprocessing. Parameterized complexity considers a problem in a two-dimensional setting, where in addition to the input size , a problem parameter  is taken into consideration. This parameter can encode a structural aspect of the problem instance. A problem is called fixed-parameter tractable (FPT) if it can be solved in time where is a function of the parameter and is a polynomial of the input size . Thus, for FPT problems, the combinatorial explosion can be confined to the parameter and is independent of the input size. It is known that a problem is fixed-parameter tractable if and only if every problem input can be reduced by polynomial-time preprocessing to an equivalent input whose size is bounded by a function of the parameter [Downey, Fellows, and Stege1999]. The reduced instance is called the problem kernel, the preprocessing is called kernelization. The power of polynomial-time preprocessing can now be benchmarked in terms of the size of the kernel. Once a small kernel is obtained, we can apply any method of choice to solve the kernel: brute-force search, heuristics, approximation, etc. [Guo and Niedermeier2007]. Because of this flexibility a small kernel is generally preferable to a less flexible branching-based fixed-parameter algorithm. Thus, small kernels provide an additional value that goes beyond bare fixed-parameter tractability.

In general the size of the kernel is exponential in the parameter, but many important NP-hard optimization problems such as Minimum Vertex Cover, parameterized by solution size, admit polynomial kernels, see, e.g., [Bodlaender et al.2009] for references.

In previous research several NP-hard AI problems have been shown to be fixed-parameter tractable. We list some important examples from various areas:

However, only exponential kernels are known for these fundamental AI problems. Can we hope for polynomial kernels?


Our results are throughout negative. We provide strong theoretical evidence that none of the above fixed-parameter tractable AI problems admits a polynomial kernel. More specifically, we show that a polynomial kernel for any of these problems causes a collapse of the Polynomial Hierarchy to its third level, which is considered highly unlikely by complexity theorists.

Our results are general: The kernel lower bounds are not limited to a particular preprocessing technique but apply to any clever technique that could be conceived in future research. Hence the results contribute to the foundations of AI.

Our results suggest the investigation of alternative approaches to polynomial-time preprocessing; for instance, preprocessing that produces in polynomial time a Boolean combination of polynomially sized kernels instead of one single kernel.

2 Formal Background

A parameterized problem P is a subset of for some finite alphabet . For a problem instance we call the main part and the parameter. We assume the parameter is represented in unary. For the parameterized problems considered in this paper, the parameter is a function of the main part, i.e., for a function . We then denote the problem as , e.g., -CSP(width) denotes the problem -CSP parameterized by the width of the given tree decomposition.

A parameterized problem P is fixed-parameter tractable if there exists an algorithm that solves any input in time where is an arbitrary computable function of and is a polynomial in .

A kernelization for a parameterized problem is an algorithm that, given , outputs in time polynomial in a pair such that (i)  if and only if and (ii) , where is an arbitrary computable function, called the size of the kernel. In particular, for constant the kernel has constant size . If is a polynomial then we say that P admits a polynomial kernel.

Every fixed-parameter tractable problem admits a kernel. This can be seen by the following argument due to DowneyFellowsStege99 DowneyFellowsStege99. Assume we can decide instances of problem P in time . We kernelize an instance as follows. If then we already have a kernel of size . Otherwise, if , then is a polynomial; hence we can decide the instance in polynomial time and replace it with a small decision-equivalent instance . Thus we always have a kernel of size at most . However, is super-polynomial for NP-hard problems (unless ), hence this generic construction is not providing polynomial kernels.

We understand preprocessing for an NP-hard problem as a polynomial-time procedure that transforms an instance of the problem to a (possible smaller) solution-equivalent instance of the same problem. Kernelization is such a preprocessing with a performance guarantee, i.e., we are guaranteed that the preprocessing yields a kernel whose size is bounded in terms of the parameter of the given problem instance. In the literature also different forms of preprocessing have been considered. An important one is knowledge compilation, a two-phases approach to reasoning problems where in a first phase a given knowledge base is (possibly in exponential time) preprocessed (“compiled”), such that in a second phase various queries can be answered in polynomial time [Cadoli et al.2002].

3 Tools for Kernel Lower Bounds

In the sequel we will use recently developed tools to obtain kernel lower bounds. Our kernel lower bounds are subject to the widely believed complexity theoretic assumption (or equivalently, ). In other words, the tools allow us to show that a parameterized problem does not admit a polynomial kernel unless the Polynomial Hierarchy collapses to its third level (see, e.g., Papadimitriou94, Papadimitriou94).

A composition algorithm for a parameterized problem is an algorithm that receives as input a sequence , uses time polynomial in , and outputs with (i)  if and only if for some , and (ii)  is polynomial in . A parameterized problem is compositional if it has a composition algorithm. With each parameterized problem we associate a classical problem

where denotes an arbitrary symbol from and is a new symbol not in . We call the unparameterized version of P.

The following result is the basis for our kernel lower bounds.

Theorem 1 (BodlaenderDowneyFellowsHermelin09, BodlaenderDowneyFellowsHermelin09, FortnowSanthanam08, FortnowSanthanam08).

Let P be a parameterized problem whose unparameterized version is NP-complete. If P is compositional, then it does not admit a polynomial kernel unless , i.e., the Polynomial Hierarchy collapses.

Let be parameterized problems. We say that P is polynomial parameter reducible to Q if there exists a polynomial time computable function and a polynomial , such that for all we have (i)  if and only if , and (ii) . The function is called a polynomial parameter transformation.

The following theorem allows us to transform kernel lower bounds from one problem to another.

Theorem 2 (BodlaenderThomasseYeo09, BodlaenderThomasseYeo09).

Let P and Q be parameterized problems such that is NP-complete, is in NP, and there is a polynomial parameter transformation from P to Q. If Q has a polynomial kernel, then P has a polynomial kernel.

4 Constraint Networks

Constraint networks have proven successful in modeling everyday cognitive tasks such as vision, language comprehension, default reasoning, and abduction, as well as in applications such as scheduling, design, diagnosis, and temporal and spatial reasoning [Dechter2010]. A constraint network is a triple where is a finite set of variables, is a finite universe of values, and is set of constraints. Each constraint is a pair where is a list of variables of length called the constraint scope, and is an -ary relation over , called the constraint relation. The tuples of indicate the allowed combinations of simultaneous values for the variables . A solution is a mapping such that for each and , we have . A constraint network is satisfiable if it has a solution.

With a constraint network we associate its constraint graph where contains an edge between two variables if and only if they occur together in the scope of a constraint. A width tree decomposition of a graph is a pair where is a tree and is a labeling of the nodes of with sets of vertices of such that the following properties are satisfied: (i) every vertex of belongs to for some node of ; (ii) every edge of is is contained in for some node of ; (iii) For each vertex of the set of all tree nodes with induces a connected subtree of ; (iv)  holds for all tree nodes . The treewidth of is the smallest such that has a width tree decomposition. The induced width of a constraint network is the treewidth of its constraint graph [Dechter and Pearl1989]. We note in passing that the problem of finding a tree decomposition of width is NP-hard but fixed-parameter tractable in .

Let be a fixed universe containing at least two elements. We consider the following parameterized version of the constraint satisfaction problem (CSP).

Instance: A constraint network and a width tree decomposition of the constraint graph of .
Parameter: The integer .
Question: Is satisfiable?

It is well known that -CSP(width) is fixed-parameter tractable over any fixed universe  [Dechter and Pearl1989, Gottlob, Scarcello, and Sideri2002] (for generalizations see SamerSzeider10a, SamerSzeider10a). We contrast this classical result and show that it is unlikely that -CSP(width) admits a polynomial kernel, even in the simplest case where .

Theorem 3.

-CSP(width) does not admit a polynomial kernel unless the Polynomial Hierarchy collapses.


We show that -CSP(width) is compositional. Let , , be a given sequence of instances of -CSP(width) where is a constraint network and is a width tree decomposition of the constraint graph of . We may assume, w.l.o.g., that for (otherwise we can simply change the names of variables). We form a new constraint network as follows. We put where are new variables. We define the set of constraints in three groups.

(1) For each and each constraint we add to a new constraint where .

(2) We add ternary constraints where and , , .

(3) Finally, we add two unary constraints and which force the values of and to and , respectively.

Let be the constraint graphs of and , respectively. Fig. 1 shows an illustration of  for .

Figure 1: Constraint graph .

We observe that are cut vertices of . Removing these vertices separates into independent parts where is the path , and is isomorphic to . By standard techniques (see, e.g., Kloks94, Kloks94), we can put the given width tree decompositions of and the trivial width 1 tree decomposition of together to a width tree decomposition of . Clearly can be obtained from , , in polynomial time.

We claim that is satisfiable if and only if at least one of the is satisfiable. This claim can be verified by means of the following observations: The constraints in groups (2) and (3) provide that for any satisfying assignment there will be some such that are all set to 0 and are all set to ; consequently is set to and all for are set to 1. The constraints in group (1) provide that if we set to , then we obtain from the original constraint ; if we set to then we obtain a constraint that can be satisfied by setting all remaining variables to . We conclude that -CSP(width) is compositional.

In order to apply Theorem 1, it remains to establish that the unparameterized version of -CSP(width) is NP-complete. Deciding whether a constraint network  over the universe is satisfiable is well-known to be NP-complete (say by reducing 3-SAT). To a constraint network on variables we can always add a trivial width tree decomposition of its constraint graph (taking a single tree node where contains all variables of ). Hence is NP-complete.

5 Satisfiability

The propositional satisfiability problem (SAT) was the first problem shown to be NP-hard [Cook1971]. Despite its hardness, SAT solvers are increasingly leaving their mark as a general-purpose tool in areas as diverse as software and hardware verification, automatic test pattern generation, planning, scheduling, and even challenging problems from algebra [Gomes et al.2008]. SAT solvers are capable of exploiting the hidden structure present in real-world problem instances. The concept of backdoors, introduced by WilliamsGomesSelman03a WilliamsGomesSelman03a provides a means for making the vague notion of a hidden structure explicit. Backdoors are defined with respect to a “sub-solver” which is a polynomial-time algorithm that correctly decides the satisfiability for a class of CNF formulas. More specifically, GomesKautzSabharwalSelman08 GomesKautzSabharwalSelman08 define a sub-solver to be an algorithm that takes as input a CNF formula and has the following properties: (i) Trichotomy: either rejects the input , or determines correctly as unsatisfiable or satisfiable; (ii) Efficiency: runs in polynomial time; (iii) Trivial Solvability: can determine if is trivially satisfiable (has no clauses) or trivially unsatisfiable (contains only the empty clause); (iv.) Self-Reducibility: if determines , then for any variable and value , determines . denotes the formula obtained from by applying the partial assignment , i.e., satisfied clauses are removed and false literals are removed from the remaining clauses.

We identify a sub-solver with the class of CNF formulas whose satisfiability can be determined by . A strong -backdoor set (or -backdoor, for short) of a CNF formula is a set of variables such that for each possible truth assignment to the variables in , the satisfiability of can be determined by sub-solver in time . Hence, if we know an -backdoor of size , we can decide the satisfiability of by running on instances , yielding a time bound of . Hence SAT decision is fixed-parameter tractable in the backdoor size  for any sub-solver . Hence the following problem is clearly fixed-parameter tractable for any sub-solver .

Instance: A CNF formula , and an -backdoor of of size .
Parameter: The integer .
Question: Is satisfiable?

We are concerned with the question of whether instead of trying all possible partial assignments we can reduce the instance to a polynomial kernel. We will establish a very general result that applies to all possible sub-solvers.

Theorem 4.

SAT(-backdoor) does not admit a polynomial kernel for any sub-solver unless the Polynomial Hierarchy collapses.


We will devise polynomial parameter transformations from the following parameterized problem which is known to be compositional [Fortnow and Santhanam2008] and therefore unlikely to admit a polynomial kernel.

Instance: A propositional formula in CNF on variables.

Parameter: The number of variables.
Question: Is satisfiable?

Let be a CNF formula and the set of all variables of . Due to property (ii) of a sub-solver, is an -backdoor set for any . Hence, by mapping (as an instance of SAT(vars)) to (as an instance of SAT(-backdoor)) provides a (trivial) polynomial parameter transformation from SAT(vars) to SAT(-backdoor). Since the unparameterized versions of both problems are clearly NP-complete, the result follows by Theorem 2. ∎

Let 3SAT() (where is an arbitrary parameterization) denote the problem SAT() restricted to 3CNF formula, i.e., to CNF formulas where each clauses contains at most three literals. In contrast to SAT(vars), the parameterized problem 3SAT(vars) has a trivial polynomial kernel: if we remove duplicate clauses, then any 3CNF formula on variables contains at most clauses, and so is a polynomial kernel. Hence the easy proof of Theorem 4 does not carry over to 3SAT(-backdoor). We therefore consider the cases 3SAT(Horn-backdoor) and 3SAT(2CNF-backdoor) separately, these cases are important since the detection of Horn and 2CNF-backdoors is fixed-parameter tractable [Nishimura, Ragde, and Szeider2004].

Theorem 5.

Neither 3SAT(Horn-backdoor) nor 3SAT(2CNF-backdoor) admit a polynomial kernel unless the Polynomial Hierarchy collapses.


Let . We show that 3SAT(-backdoor) is compositional. Let , , be a given sequence of instances of 3SAT(-backdoor) where is a 3CNF formula and is a -backdoor set of of size . We distinguish two cases.

Case 1: . Let and . Whether is satisfiable or not can be decided in time since the satisfiability of a Horn or 2CNF formula can be decided in linear time. We can check whether at least one of the formulas is satisfiable in time which is polynomial in . If some is satisfiable, we output ; otherwise we output ( is unsatisfiable). Hence we have a composition algorithm.

Case 2: . This case is more involved. We construct a new instance of 3SAT(-backdoor) as follows.

Let . Since , follows.

Let denote the set of variables of . We may assume, w.l.o.g., that and that for all since otherwise we can change names of variable accordingly. In a first step we obtain from every a CNF formula as follows. For each variable we take two new variables and . We replace each positive occurrence of a variable in with the literal and each negative occurrence of with the literal . We add all clauses of the form for ; we call these clauses “connection clauses.” Let be the formula obtained from in this way. We observe that and are SAT-equivalent, since the connection clauses form an implication chain. Since the connection clauses are both Horn and 2CNF, is also a -backdoor of .

We take a set of new variables. Let be the sequence of all possible clauses (modulo permutation of literals within a clause) containing exactly literals over the variables in . Consequently we can write as where .

For we add to each connection clause of the literal . Let denote the 3CNF formula obtained from this way.

For we define 3CNF formulas as follows. If then consists just of the clause . If then we take new variables and let consist of the clauses , , . Finally, we let be the 3CNF formula containing all the clauses from . Any assignment to that satisfies can be extended to an assignment that satisfies since such assignment satisfies at least one connection clause and so the chain of implications from from to is broken.

It is not difficult to verify the following two claims. (i)  is satisfiable if and only if at least one of the formulas is satisfiable. (ii)  is a -backdoor of . Hence we have also a composition algorithm in Case 2, and thus 3SAT(-backdoor) is compositional. Clearly is NP-complete, hence the result follows from Theorem 1. ∎

6 Global Constraints

The success of today’s constraint solvers relies heavily on efficient algorithms for special purpose global constraints [van Hoeve and Katriel2006]. A global constraint specifies a pattern that frequently occurs in real-world problems, for instance, it is often required that variables must all take different values (e.g., activities requiring the same resource must all be assigned different times). The AllDifferent global constraint efficiently encodes this requirement.

More formally, a global constraint is defined for a set of variables, each variable ranges over a finite domain of values. An instantiation is an assignment such that for each . A global constraint defines which instantiations are legal and which are not. A global constraint is consistent if it has at least one legal instantiation, and it is domain consistent (or hyper arc consistent) if for each variable and each value there is a legal instantiation with . For all global constraints considered in this paper, domain consistency can be reduced to a quadratic number of consistency checks, hence we will focus on consistency. We assume that the size of a representation of a global constraint is polynomial in .

For several important types of global constraints, the problem of deciding whether a constraint of type is consistent (in symbols -Cons) is NP-hard. Examples for such intractable types of constraints are NValue, Disjoint, and Uses [Bessière et al.2004]. An NValue constraint over a set of variables requires from a legal instantiation that ; AllDifferent is the special case where . The global constraints Disjoint and Uses are specified by two sets of variables ; Disjoint requires that for each pair and ; Uses requires that for each there is some such that . For a set of variables we write .

BessiereEtal08 BessiereEtal08 considered as parameter for NValue, as parameter for Disjoint, and as parameter for Uses. They showed that consistency checking is fixed-parameter tractable for the constraints under the respective parameterizations, i.e., the problems NValue-Cons(), Disjoint-Cons(), and Uses-Cons() are fixed-parameter tractable. We show that it is unlikely that their results can be improved in terms of polynomial kernels.

Theorem 6.

The problems NValue-Cons(), Disjoint-Cons(), Uses-Cons() do not admit polynomial kernels unless the Polynomial Hierarchy collapses.


We devise a polynomial parameter reduction from SAT(vars). We use a construction of BessiereEtAl04 BessiereEtAl04. Let be a CNF formula over variables . We consider the clauses and variables of as the variables of a global constraint with domains , and . Now can be encoded as an NValue constraint with and (clearly is satisfiable if and only if the constraint is consistent). Since we have a polynomial parameter reduction from SAT(vars) to NValue-Cons(). Similarly, as observed by BessiereHebrardHnichKiziltanWalsh09 BessiereHebrardHnichKiziltanWalsh09, can be encoded as a Disjoint constraint with and (), or as a Uses constraint with and (). Since the unparameterized problems are clearly NP-complete, the result follows by Theorem 2. ∎

Further results on kernels for global constraints have been obtained by GaspersSzeider11 GaspersSzeider11.

7 Bayesian Reasoning

Bayesian networks (BNs) have emerged as a general representation scheme for uncertain knowledge [Pearl2010]

. A BN models a set of stochastic variables, the independencies among these variables, and a joint probability distribution over these variables. For simplicity we consider the important special case where the stochastic variables are Boolean. The variables and independencies are modelled in the BN by a directed acyclic graph

, the joint probability distribution is given by a table for each node which defines a probability for each possible instantiation of the parents of in . The probability of a complete instantiation of the variables of  is given by the product of over all variables . We consider the problem Positive-BN-Inference which takes as input a Boolean BN and a variable , and asks whether . The problem is NP-complete [Cooper1990] and moves from NP to #P if we ask to compute [Roth1996]. The problem can be solved in polynomial time if the BN is singly connected, i.e, if there is at most one undirected path between any two variables [Pearl1988]. It is natural to parametrize the problem by the number of variables one must delete in order to make the BN singly connected (the deleted variables form a loop cutset). In fact, Positive-BN-Inference(loop cutset size) is easily seen to be fixed-parameter tractable as we can determine whether by taking the maximum of over all possible instantiations of the cutset variables, each of which requires processing of a singly connected network. However, although fixed-parameter tractable, it is unlikely that the problem admits a polynomial kernel.

Theorem 7.

Positive-BN-Inference(loop cutset size) does not admit a polynomial kernel unless the Polynomial Hierarchy collapses.


(Sketch.) We give a polynomial parameter transformation from SAT(vars) and apply Theorem 2. The reduction is based on the reduction from 3SAT given by Cooper90 Cooper90. However, we need to allow clauses with an arbitrary number of literals since, as observed above, 3SAT(vars) has a polynomial kernel. Let be a CNF formula on variables. We construct a BN such that for a variable we have if and only if is satisfiable. Cooper uses input nodes for representing variables of , clause nodes for representing the clauses of , and conjunction nodes for representing the conjunction of the clauses. We proceed similarly, however, we cannot represent a clause of large size with a single clause node , as the required table would be of exponential size. Therefore we split clauses containing more than 3 literals into several clause nodes, as indicated in Figure 2.

Figure 2: BN representation of a clause on four literals.

It remains to observe that the set of input nodes is a loop cutset of the constructed BN, hence we have indeed a polynomial parameter transformation from SAT(vars) to Positive-BN-Inference(loop cutset size). The result follows by Theorem 2. ∎

8 Nonmonotonic Reasoning

Logic programming with negation under the stable model semantics is a well-studied form of nonmonotonic reasoning [Gelfond and Lifschitz1988, Marek and Truszczyński1999]. A (normal) logic program  is a finite set of rules of the form

where are atoms, where forms the head and the from the body of . We write , , and . Let  be a finite set of atoms. The GF reduct of a logic program  under is the program obtained from  by removing all rules with , and removing from the body of each remaining rule all literals with . is a stable model of if is a minimal model of , i.e., if (i) for each rule with we have , and (ii) there is no proper subset of with this property. The undirected dependency graph of is formed as follows. We take the atoms of as vertices and add an edge between two atoms if there is a rule with and , and we add a path if and ( is a new vertex of degree 2). The feedback width of is the size of a smallest set of atoms such that every cycle of runs through an atom in .

A fundamental computational problems is Stable Model Existence (SME), which asks whether a given normal logic program has a stable model. The problem is well-known to be NP-complete [Marek and Truszczyński1991]. GottlobScarcelloSideri02 GottlobScarcelloSideri02 showed that SME(feedback width) is fixed-parameter tractable (see FichteSzeider11 FichteSzeider11 for generalizations). We show that this result cannot be strengthened with respect to a polynomial kernel.

Theorem 8.

SME(feedback width) does not admit a polynomial kernel unless the Polynomial Hierarchy collapses.


(Sketch.) We give a polynomial parameter transformation from SAT(vars) to feedback width using a construction of Niemela99 Niemela99. Given a CNF formula on variables, we construct a logic program as follows. For each variable of we take two atoms and and include the rules and ; for each clause of we take an atom and include for each positive literal of the rule , and for each negative literal of the rule ; finally, we take two atoms and and include the rule and for each clause of the rule . is satisfiable if and only if has a stable model [Niemelä1999]. It remains to observe that each cycle of runs through a vertex in , hence the feedback width of is at most . Hence we have a polynomial parameter transformation from SAT(vars) to SME(feedback width). The result follows by Theorem 2. ∎

9 Conclusion

We have established super-polynomial kernel lower bounds for a wide range of important AI problems, providing firm limitations for the power of polynomial-time preprocessing for these problems. We conclude from these results that in contrast to many optimization problems (see Section 1), typical AI problems do not admit polynomial kernels. Our results suggest the consideration of alternative approaches. For example, it might still be possible that some of the considered problems admit polynomially sized Turing kernels, i.e., a polynomial-time preprocessing to a Boolean combination of a polynomial number of polynomial kernels. In the area of optimization, parameterized problems are known that do not admit polynomial kernels but admit polynomial Turing kernels [Fernau et al.2009]. This suggests a theoretical and empirical study of Turing kernels for the AI problems considered.


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