** Counting Problems **

Today we describe counting problems and the class that they define, and we show that every counting problem can be approximately solved in randomized polynomial given access to an oracle.

**1. Counting Classes **

*[Note: we did not cover most of the material of this section in class.]*

Recall that is an * NP-relation*, if there is a polynomial time algorithm such that and there is a polynomial such that .

Definition 1If is an relation, then is the problem that, given , asks how many satisfy .

is the class of all problems of the form , where is anNP-relation.

Observe that an **NP**-relation naturally defines an language , where , and every language can be defined in this way. Therefore problems in can always be seen as the problem of counting the number of witnesses for a given instance of an problem.

Unlike for decision problems there is no canonical way to define reductions for counting classes. There are two common definitions.

Definition 2We say there is aparsimonious reductionfrom to (written ) if there is a polynomial time transformation such that for all , .

Often this definition is a little too restrictive and we use the following definition instead.

Definition 3if there is a polynomial time algorithm for given an oracle that solves .

is the problem where given a circuit, we want to count the number of inputs that make the circuit output 1.

Theorem 4is -complete under parsimonious reductions.

*Proof:* Let be in and and be as in the definition. Given we want to construct a circuit such that . We then construct that on input simulates . From earlier arguments we know that this can be done with a circuit with size about the square of the running time of . Thus will have size polynomial in the running time of and so polynomial in . Then let .

Theorem 5is -complete.

*Proof:* We show that there is a parsimonious reduction from to . That is, given a circuit we construct a Boolean formula such that the number of satisfying assignments for is equal to the number of inputs for which outputs 1. Suppose has inputs and gates and has inputs , where the represent the output of gate . Now each gate has two input variables and one output variable. Thus a gate can be complete described by mimicking the output for each of the 4 possible inputs. Thus each gate can be simulated using at most 4 clauses. In this way we have reduced to a formula with variables and clauses. So there is a parsimonious reduction from to .

Notice that if a counting problem is -complete under parsimonious reductions, then the associated language is **NP**-complete, because implies . On the other hand, with the less restrictive definition of reducibility, even some counting problems whose decision version is in **P** are -complete. For example, the problem of counting the number of satisfying assignments for a given 2CNF formula and the problem of counting the number of perfect matchings in a given bipartite graphs are both -complete.

**2. Complexity of counting problems **

We will prove the following theorem:

Theorem 6For every counting problem in , there is a probabilistic algorithm that on input , computes with high probability a value such that

in time polynomial in and in , using an oracle for .

The theorem says that can be approximate in . We remark that approximating is -hard, and so to compute an approximation we need at least the power of . Theorem 6 states that the power of and randomization is sufficient.

Another remark concerns the following result.

Theorem 7 (Toda)For every , .

This implies that is -hard for every , i.e., lies outside the polynomial hierarchy, unless the hierarchy collapses. Recall that lies inside , and hence approximating can be done in . Therefore, approximating cannot be equivalent to computing exactly, unless the polynomial hierarchy collapses.\footnote{The above discussion was not very rigorous but it can be correctly formalized. In particular: (i) from the fact that and that approximate counting is doable in it does not necessarily follow that approximate counting is in , although in this case it does because the proof that relativizes; (ii) we have defined , , etc., as classes of decision problems, while approximate counting is not a decision problem (it can be shown, however, to be equivalent to a “promise problem,” and the inclusion holds also for promise problems.}

We first make some observations so that we can reduce the proof to the task of proving a simpler statement.

- It is enough to prove the theorem for .
If we have an approximation algorithm for , we can extend it to any in using the parsimonious reduction from to .

- It is enough to give a polynomial time -approximation for .
Suppose we have an algorithm and a constant such that

Given , we can construct where each is a copy of constructed using fresh variables. If has satisfying assignments, has satisfying assignments. Then, giving to the algorithm we get

If is a constant and , .

- For a formula that has satisfying assignments, can be found in .
This can be done by iteratively asking the oracle the questions of the form: “Are there assignments satisfying this formula?” Notice that these are questions, because the algorithm can guess these assignments and check them.

**3. An approximate comparison procedure **

Suppose that we had available an approximate comparison procedure * a-comp* with the following properties:

- If then with high probability;
- If then with high probability.

Given * a-comp*, we can construct an algorithm that -approximates as described below:

- Input:
- compute:
- if outputs NO from the first time then
- //
*The value is either or and the answer can be checked by one more query to the oracle.* - Query to the oracle and output an exact value.

- //
- else
- Suppose that it outputs YES for and NO for
- Output

We need to show that this algorithm approximates within a factor of . If * a-comp* answers NO from the first time, the algorithm outputs the right answer because it checks for the answer explicitly. Now suppose * a-comp* says YES for all and says NO for . Since outputs YES, , and also since outputs NO, . The algorithm outputs . Hence,

and the algorithm outputs the correct answer with in a factor of .

Thus, to establish the theorem, it is enough to give a implementation of the * a-comp* procedure

**4. Constructing a-comp **

The procedure and its analysis is similar to the Valiant-Vazirani reduction: for a given formula we pick a hash function from a pairwise independent family, and look at the number of assignments that satisfy and such that .

In the Valiant-Vazirani reduction, we proved that if is a set of size approximately equal to the size of the range of , then, with constant probability, exactly one element of is mapped by into . Now we use a different result, a simplified version of the “Leftover Hash Lemma” proved by Impagliazzo, Levin, and Luby in 1989, that says that if is sufficiently larger than the range of then the number of elements of mapped into is concentrated around its expectation.

Lemma 8Let be a family of pairwise independent hash functions . Let , . Then,

From this, * a-comp* can be constructed as follows.

- input:
- if then check exactly whether .
- if
- pick from a set of pairwise independent hash functions , where
- answer YES iff there are more then assignments to such that satisfies and .

Notice that the test at the last step can be done with one access to an oracle to . We will show that the algorithm is in . Let be the set of satisfying assignments for . There are 2 cases.

- If , by Lemma 8 we have:
(set , and , because )

which is the success probability of the algorithm.

- If :
Let be a superset of of size . We have

(by Lemma 8 with .)

Therefore, the algorithm will give the correct answer with probability at least , which can then be amplified to, say, (so that all invocations of * a-comp* are likely to be correct) by repeating the procedure times and taking the majority answer.

**5. The Remaining Proof **

We finish the lecture by proving Lemma 8.

*Proof:* We will use Chebyshev’s Inequality to bound the failure probability. Let , and pick a random . We define random variables as

Clearly, .

We now calculate the expectations. For each , and . Hence,

Also we calculate the variance

Because are pairwise independent,

Using Chebyshev’s Inequality, we get

**6. Approximate Sampling **

*The content of this section was not covered in class; it’s here as bonus material. It’s good stuff.*

So far we have considered the following question: for an **NP**-relation , given an input , what is the size of the set ? A related question is to be able to sample from the uniform distribution over .

Whenever the relation is “downward self reducible” (a technical condition that we won’t define formally), it is possible to prove that there is a probabilistic algorithm running in time polynomial in and to approximate within the value if and only if there is a probabilistic algorithm running in time polynomial in and that samples a distribution -close to the uniform distribution over .

We show how the above result applies to 3SAT (the general result uses the same proof idea). For a formula , a variable and a bit , let us define by the formula obtained by substituting the value in place of .\footnote{ Specifically, is obtained by removing each occurrence of from the clauses where it occurs, and removing all the clauses that contain an occurrence of ; the formula is similarly obtained.}

If is defined over variables , it is easy to see that

Also, if is the uniform distribution over satisfying assignments for , we note that

Suppose then that we have an efficient sampling algorithm that given and generates a distribution -close to uniform over the satisfying assignments of .

Let us then ran the sampling algorithm with approximation parameter and use it to sample about assignments. By computing the fraction of such assignments having and , we get approximate values , such that . Let be such that , then is a good approximation, to within a multiplicative factor to , and we can recurse to compute to within a factor.

Conversely, suppose we have an approximate counting procedure. Then we can approximately compute , generate a value for with probability approximately , and then recurse to generate a random assignment for .

The same equivalence holds, clearly, for 2SAT and, among other problems, for the problem of counting the number of perfect matchings in a bipartite graph. It is known that it is **NP**-hard to perform approximate counting for 2SAT and this result, with the above reduction, implies that approximate sampling is also hard for 2SAT. The problem of approximately sampling a perfect matching has a probabilistic polynomial solution, and the reduction implies that approximately counting the number of perfect matchings in a graph can also be done in probabilistic polynomial time.

The reduction and the results from last section also imply that 3SAT (and any other relation) has an approximate sampling algorithm that runs in probabilistic polynomial time with an oracle. With a careful use of the techniques from last week it is indeed possible to get an *exact* sampling algorithm for 3SAT (and any other relation) running in probabilistic polynomial time with an oracle. This is essentially best possible, because the approximate sampling requires randomness by its very definition, and generating satisfying assignments for a 3SAT formula requires at least an oracle.

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December 1, 2011 at 9:54 pm

anyeOh is there no further lecture notes?