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- Introduce the class \(\mathbf{NP}\) capturing a great many important
computational problems
- \(\mathbf{NP}\)-completeness: evidence that a problem might be
intractable.
- The \(\mathbf{P}\) vs \(\mathbf{NP}\) problem.

“In this paper we give theorems that suggest, but do not imply, that these problems, as well as many others, will remain intractable perpetually”, Richard Karp, 1972

So far we have shown that 3SAT is no harder than Quadratic Equations,
Independent Set, Maximum Cut, and Longest Path. But to show that these
problems are *computationally equivalent* we need to give reductions in
the other direction, reducing each one of these problems to 3SAT as
well. It turns out we can reduce all three problems to 3SAT in one fell
swoop.

In fact, this result extends far beyond these particular problems. All
of the problems we discussed in reductionchap, and a great
many other problems, share the same commonality: they are all *search*
problems, where the goal is to decide, given an instance \(x\), whether
there exists a *solution* \(y\) that satisfies some condition that can be
verified in polynomial time. For example, in 3SAT, the instance is a
formula and the solution is an assignment to the variable; in Max-Cut
the instance is a graph and the solution is a cut in the graph; and so
on and so forth. It turns out that *every* such search problem can be
reduced to 3SAT.

To make this precise, we make the following mathematical definition: we
define the class \(\mathbf{NP}\) to contain all Boolean functions that
correspond to a *search problem* of the form above\(-\) that is, functions
that output \(1\) on \(x\) if and only if there exists a solution \(w\) such
that the pair \((x,w)\) satisfies some polynomial-time checkable
condition. Formally, \(\mathbf{NP}\) is defined as follows:

We say that \(F:\{0,1\}^* \rightarrow \{0,1\}\) is in \(\mathbf{NP}\) if there exists some constants \(a,b \in \N\) and \(V:\{0,1\}^* \rightarrow \{0,1\}\) such that \(V\in \mathbf{P}\) and for every \(x\in \{0,1\}^n\) \[ F(x)=1 \Leftrightarrow \exists_{w \in \{0,1\}^{an^b}} \text{ s.t. } V(xw)=1 \label{NP:eq} \]

See also NPdeffigfig for an illustration of NP-def.
The name \(\mathbf{NP}\) stands for “nondeterministic polynomial time” and
is used for historical reasons; see the bibiographical notes. The string
\(w\) in \eqref{{NP:eq}} is sometimes known as a *solution*,
*certificate*, or *witness* for the instance \(x\).

The definition of \(\mathbf{NP}\) means that for every \(F\in \mathbf{NP}\)
and string \(x\in \{0,1\}^*\), \(F(x)=1\) if and only if there is a *short
and efficiently verifiable proof* of this fact. That is, we can think of
the function \(G\) in NP-def as a *verifier* algorithm, similar
to what we’ve seen in proofdef. The verifier checks whether a
given string \(w\in \{0,1\}^*\) is a valid proof for the statement
“\(F(x)=1\)”. Essentially all proof systems considered in mathematics
involve line-by-line checks that can be carried out in polynomial time.
Thus the heart of \(\mathbf{NP}\) is asking for statements that have
*short* (i.e., polynomial in the size of the statements) proof. As we
will see later on, for this reason Kurt Gödel phrased the question of
whether \(\mathbf{NP}=\mathbf{P}\) as asking whether “the mental work of a
mathematician [in proving theorems] could be completely replaced by a
machine”.

The NP-def is *assymetric* in the sense that there is a
difference between an output of \(1\) and an output of \(0\). You should
make sure you understand why this definition does *not* guarantee that
if \(F \in \mathbf{NP}\) then the function \(1-F\) is in \(\mathbf{NP}\) as
well. In fact, this is believed *not* to be the case in general. This is
in contrast to the class \(\mathbf{P}\) which (as you should verify)
*does* satisfy that if \(F\in \mathbf{P}\) then \(1-F\) is in \(\mathbf{P}\)
as well.

- \(3SAT\) is in \(\mathbf{NP}\) since for every \(\ell\)-variable formula
\(\varphi\), \(3SAT(\varphi)=1\) if and only if there exists a
satisfying assignment \(x \in \{0,1\}^\ell\) such that \(\varphi(x)=1\),
and we can check this condition in polynomial time.
Note that an \(\ell\) variable formula \(\varphi\) is represented by a string of length at least \(\ell\), and we can use some “padding” in our encoding so that the assignment to \(\varphi\)’s variables is encoded by a string of length exactly \(|\varphi|\). We can always use this padding trick, and so one can think of the condition \eqref{{NP:eq}} as simply stipulating that the “solution” \(y\) to the problem \(x\) is of size at most \(poly(|x|)\). - \(QUADEQ\) is in \(\mathbf{NP}\) since for every \(\ell\)-variable instance of quadratic equations \(E\), \(QUADEQ(E)=1\) if and only if there exists an assignment \(x\in \{0,1\}^\ell\) that satisfies \(E\), and we can check this condition in polynomial time.
- \(ISET\) is in \(\mathbf{NP}\) since for every graph \(G\) and integer \(k\), \(ISET(G,k)=1\) if and only if there exists a set \(S\) of \(k\) vertices that contains no pair of neighbors in \(G\), and we can check this condition in polynomial time.
- \(LONGPATH\) is in \(\mathbf{NP}\) since for every graph \(G\) and integer \(k\), \(LONGPATH(G,k)=1\) if and only if there exists a simple path \(P\) in \(G\) that is of length at least \(k\), and we can check this condition in polynomial time.
- \(MAXCUT\) is in \(\mathbf{NP}\) since for every graph \(G\) and integer \(k\), \(MAXCUT(G,k)=1\) if and only if there exists a cut \((S,\overline{S})\) in \(G\) that cuts at least \(k\) edges, and we can check this condition in polynomial time.

There are many, many, *many*, more examples of interesting functions we
would like to compute that are easily shown to be in \(\mathbf{NP}\). What
is quite amazing is that if we can solve 3SAT then we can solve all of
them!

The following is one of the most fundamental theorems in Computer Science:

For every \(F\in \mathbf{NP}\), \(F \leq_p 3SAT\).

We will soon show the proof of cook-levin-thm, but note that
it immediately implies that \(QUADEQ\), \(LONGPATH\), and \(MAXCUT\) all
reduce to \(3SAT\). In fact, combining it with the reductions we’ve seen,
it implies that all these problems are *equivalent!* For example, to
reduce \(QUADEQ\) to \(LONGPATH\), we can first reduce \(QUADEQ\) to \(3SAT\)
using cook-levin-thm and use the reduction we’ve seen from
\(3SAT\) to \(LONGPATH\). There is of course nothing special about \(QUADEQ\)
here\(-\) by combining \eqref{cook-levin-thm} with the reduction we
saw, we see that just like \(3SAT\), *every* \(F\in \mathbf{NP}\) reduces to
\(LONGPATH\), and the same is true for \(QUADEQ\) and \(MAXCUT\). All these
problems are in some sense “the hardest in \(\mathbf{NP}\)” since an
efficient algorithm for any one of them would imply an efficient
algorithm for *all* the problems in \(\mathbf{NP}\). This motivates the
following definition:

We say that \(G:\{0,1\}^* \rightarrow \{0,1\}\) is *\(\mathbf{NP}\) hard* if
for every \(F\in \mathbf{NP}\), \(F \leq_p G\).

We say that \(G:\{0,1\}^* \rightarrow \{0,1\}\) is *\(\mathbf{NP}\)
complete* if \(G\) is \(\mathbf{NP}\) hard and \(G\) is in \(\mathbf{NP}\).

cook-levin-thm and the reductions we’ve seen in the last
lecture show that despite their superficial differences, 3SAT, quadratic
equations, longest path, independent set, and maximum cut, are all
\(\mathbf{NP}\)-complete. Many thousands of additional problems have been
shown to be \(\mathbf{NP}\)-complete, arising from all the sciences,
mathematics, economics, engineering and many other fields.

Clearly \(\mathbf{NP} \supseteq \mathbf{P}\), since if we can decide efficiently whether \(F(x)=1\), we can simply ignore any “solution” that we are presented with. (However, it is still an excellent idea for you to pause here and verify that you see why every \(F\in \mathbf{P}\) will be in \(\mathbf{NP}\) as per NP-def.) Also, \(\mathbf{NP} \subseteq \mathbf{EXP}\), since all the problems in \(\mathbf{NP}\) can be solved in exponential time by enumerating all the possible solutions. (Again, please verify that you understand why this follows from the definition.)

*The* most famous conjecture in Computer Science is that
\(\mathbf{P} \neq \mathbf{NP}\). One way to refute this conjecture is to
give a polynomial-time algorithm for even a single one of the
\(\mathbf{NP}\)-complete problems such as 3SAT, Max Cut, or the thousands
of others that have been studied in all fields of human endeavors. The
fact that these problems have been studied by so many people, and yet
not a single polynomial-time algorithm for any of them has been found,
supports that conjecture that indeed \(\mathbf{P} \neq \mathbf{NP}\). In
fact, for many of these problems (including all the ones we mentioned
above), we don’t even know of a \(2^{o(n)}\)-time algorithm! However, to
the frustration of computer scientists, we have not yet been able to
prove that \(\mathbf{P}\neq\mathbf{NP}\) or even rule out the existence of
an \(O(n)\)-time algorithm for 3SAT. Resolving whether or not
\(\mathbf{P}=\mathbf{NP}\) is known as the \(\mathbf{P}\) vs \(\mathbf{NP}\)
problem. A
million-dollar prize has been
offered
for the solution of this problem, a popular
book has been written, and every
year a new paper comes out claiming a proof of \(\mathbf{P}=\mathbf{NP}\)
or \(\mathbf{P}\neq\mathbf{NP}\), only to wither under scrutiny.

One of the mysteries of computation is that people have observed a certain empirical “zero-one law” or “dichotomy” in the computational complexity of natural problems, in the sense that many natural problems are either in \(\mathbf{P}\) (often in \(TIME(O(n))\) or \(TIME(O(n^2))\)), or they are are \(\mathbf{NP}\) hard. This is related to the fact that for most natural problems, the best known algorithm is either exponential or polynomial, with not too many examples where the best running time is some strange intermediate complexity such as \(2^{2^{\sqrt{\log n}}}\). However, it is believed that there exist problems in \(\mathbf{NP}\) that are neither in \(\mathbf{P}\) nor are \(\mathbf{NP}\)-complete, and in fact a result known as “Ladner’s Theorem” shows that if \(\mathbf{P} \neq \mathbf{NP}\) then this is indeed the case (see also ladner-ex and PNPscenariosfig).

We will now prove the Cook-Levin Theorem, which is the underpinning to a great web of reductions from 3SAT to thousands of problems across great many fields. Some problems that have been shown to be \(\mathbf{NP}\)-complete include: minimum-energy protein folding, minimum surface-area foam configuration, map coloring, optimal Nash equilibrium, quantum state entanglement, minimum supersequence of a genome, minimum codeword problem, shortest vector in a lattice, minimum genus knots, positive Diophantine equations, integer programming, and many many more. The worst-case complexity of all these problems is (up to polynomial factors) equivalent to that of 3SAT, and through the Cook-Levin Theorem, to all problems in \(\mathbf{NP}\).

To prove cook-levin-thm we need to show that \(F \leq_p 3SAT\) for every \(F\in \mathbf{NP}\). We will do so in three stages. We define two intermediate problems: \(NANDSAT\) and \(3NAND\). We will shortly show the definitions of these two problems, but cook-levin-thm will follow from the following three lemmas:

\(NANDSAT\) is \(\mathbf{NP}\)-hard.

\(NANDSAT \leq_p 3NAND\).

\(3NAND \leq_p 3SAT\).

From the transitivity of reductions, nand-thm, threenand-thm, and threenand-sat-thm together immediately imply that \(3SAT\) is \(\mathbf{NP}\)-hard, hence establishing cook-levin-thm. (Can you see why?) We now prove these three lemmas one by one, providing the requisite definitions as we go along.

We define the \(NANDSAT\) problem as follows. On input a string
\(Q\in \{0,1\}^*\), we define \(NANDSAT(Q)=1\) if and only if \(Q\) is a valid
representation of an \(n\)-input and single-output NAND program and there
exists some \(w\in \{0,1\}^n\) such that \(Q(w)=1\). While we don’t need
this to prove nand-thm, note that \(NANDSAT\) is in
\(\mathbf{NP}\) since we can verify that \(Q(w)=1\) using the
polyonmial-time algorithm for evaluating NAND programs.

To prove nand-thm we need to show that for every \(F\in \mathbf{NP}\), \(F \leq_p NANDSAT\). The high-level idea is that by the definition of \(\mathbf{NP}\), there is some NAND++ program \(P^*\) and some polynomial \(T(\cdot)\) such that \(F(x)=1\) if and only if there exists some \(w\) such that \(P^*(xw)\) outputs \(1\) within \(T(|x|)\) steps. Now by “unrolling the loop” of the NAND++ program \(P^*\) we can convert it into a NAND program \(Q\) that on input \(w\) will simulate \(P^*(xw)\) for \(T(|x|)\) steps. We will then get that \(NANDSAT(Q)=1\) if and only if \(F(x)=1\).

We now present the details. Let \(F \in \mathbf{NP}\). By NP-def there exists \(G \in \mathbf{P}\) and \(a,b \in \N\) such that for every \(x\in \{0,1\}^*\), \(F(x)=1\) if and only if there exists \(w\in \{0,1\}^{a|x|^b}\) such that \(G(xw)=1\). Since \(G\in \mathbf{P}\) there is some NAND++ program \(P^*\) that computes \(G\) in at most \({n'}^c\) time for some constant \(c\) where \(n'\) is the size of its input. Moreover, as shown in simpleNANDthm, we can assume without loss of generality that \(P^*\) is simple in the sense of simpleNANDpp.

To prove nand-thm we need to give a polynomial-time computable
map of every \(x^* \in \{0,1\}^*\) to a NAND program \(Q\) such that
\(F(x^*)=NANDSAT(Q)\). Let \(x^*\in \{0,1\}^*\) be such a string and let
\(n=|x^*|\) be its length. In time polynomial in \(n\), we can obtain a NAND
program \(Q^*\) of \(n+an^b\) inputs and \(|P^*|\cdot (n+an^b)^c\) lines
(where \(|P^*|\) denotes the number of lines of \(P^*\)) such that
\(Q^*(xw)=P^*(xw)\) for every \(x\in \{0,1\}^n\) and \(w\in \{0,1\}^{an^b}\).
Indeed, we can do this by simply copying and pasting \((n+an^b)^c\) times
the code of \(P^*\) one after the other, and replacing all references to
`i`

in the \(j\)-th copy with \(INDEX(j)\).`i`

index variable in
the \(j\)-th iteration. The particular formula for \(INDEX(j)\) was
given in \eqref{eqindex} but all we care is that it is
computable in time polynomial in \(j\).`validx_`

\(\expr{k}\) with `one`

if \(k<n+an^b\) and `zero`

otherwise. By
the definition of NAND++ and the fact that the original program \(P^*\)
was simple and halted within at most \((n+an^b)^c\) steps, the NAND
program \(Q^*\) agrees with \(P^*\) on every input of the form
\(xw \in \{0,1\}^{n+an^b}\).`one`

and `zero`

variables, and that assignments to
the output variable `y_0`

are “guarded” in the sense that adding
extra copies of \(P^*\) after it has already halted will not change
the output. It is not hard to ensure these properties, as shown in
simpleNANDthm.

Now we transform \(Q^*\) into \(Q\) by “hardwiring” its first \(n\) inputs to
corresond to \(x^*\). That is, we obtain a program \(Q\) such that
\(Q(w)=Q^*(x^*w)\) for every \(w\in \{0,1\}^{an^b}\) by replacing all
references to the variables `x_`

\(\expr{j}\) for \(j<n\) with either `one`

or `zero`

depending on the value of \(x^*_j\). (We also map
`x_`

\(\expr{n}\),\(\ldots\),`x_`

\(\expr{n+an^b}\) to `x_`

\(\expr{0}\), \(\ldots\),
`x_`

\(\expr{an^b-1}\) so that the number of inputs is reduced from
\(n+an^b\) to \(an^b\).) You can verify that by this construction, \(Q\) has
\(an^b\) inputs and for every \(w\in \{0,1\}^{an^b}\), \(Q(w)=Q^*(x^*w)\). We
now claim that \(NANDSAT(Q)=F(x^*)\). Indeed note that \(F(x^*)=1\) if and
only if there exists \(w\in \{0,1\}^{an^b}\) s.t. \(P^*(x^*w)=1\). But since
\(Q^*(xw)=P^*(xw)\) for every \(x,w\) of these lengths, and \(Q(w)=Q^*(x^*w)\)
it follows that this holds if and only if there exists
\(w\in \{0,1\}^{an^b}\) such that \(Q(w)=1\). But the latter condition holds
exactly when \(NANDSAT(Q)=1\).

The proof above is a little bit technical but ultimately follows quite
directly from the definition of \(\mathbf{NP}\), as well as of NAND and
NAND++ programs. If you find it confusing, try to pause here and work
out the proof yourself from these definitions, using the idea of
“unrolling the loop” of a NAND++ program. It might also be useful for
you to think how you would implement in your favorite programming
language the function `expand`

which on input a NAND++ program \(P\) and
numbers \(T,n\) would output an \(n\)-input NAND program \(Q\) of \(O(|T|)\)
lines such that for every input \(x\in \{0,1\}^n\), if \(P\) halts on \(x\)
within at most \(T\) steps and outputs \(y\), then \(Q(x)=y\).

The \(3NAND\) problem is defined as follows: the input is a logical formula \(\varphi\) on a set of variables \(z_0,\ldots,z_{r-1}\) which is an AND of constraints of the form \(z_i = NAND(z_j,z_k)\). For example, the following is a \(3NAND\) formula with \(5\) variables and \(3\) constraints:

\[ \left( z_3 = NAND(z_0,z_2) \right) \wedge \left( z_1 = NAND(z_0,z_2) \right) \wedge \left( z_4 = NAND(z_3,z_1) \right) \]

The output of \(3NAND\) on input \(\varphi\) is \(1\) if and only if there is an assignment to the variables of \(\varphi\) that makes it evaluate to “true” (that is, there is some assignment \(z \in \{0,1\}^r\) satisfying all of the constraints of \(\varphi\)). As usual, we can represent \(\varphi\) as a string, and so think of \(3NAND\) as a function mapping \(\{0,1\}^*\) to \(\{0,1\}\). We now prove threenand-thm.

To prove threenand-thm we need to give a polynomial-time map
from every NAND program \(Q\) to a 3NAND formula \(\varphi\) such that there
exists \(w\) such that \(Q(w)=1\) if and only if there exists \(z\) satisfying
\(\varphi\). This will actually follow directly from our notion of
“modification logs” or “deltas” of NAND++ programs (see
deltas). We will have a variable of \(\varphi\) corresponding to
every line of \(Q\), with a constraint ensuring that if line \(i\) has the
form `foo := bar NAND blah`

then the variable corresponding to line \(i\)
should be the NAND of the variables corresponding to the lines in which
`bar`

and `blah`

were just written to. We will also have variables
associated with the input \(w\), and use them in lines such as
`foo := x_17 NAND x_33`

or `foo := bar NAND x_55`

. Finally we add a
constraint that requires the last assignment to `y_0`

to equal \(1\). By
construction, satisfying assignments to our formula \(\varphi\) will
correspond to valid modification logs of executions of \(Q\) that end with
it outputting \(1\). Hence in particular there exists a satisfying
assignment to \(\varphi\) if and only if there is some input
\(w\in \{0,1\}^n\) on which the execution of \(Q\) on \(w\) ends in \(1\).

To prove threenand-thm we need to give a reduction from
\(NANDSAT\) to \(3NAND\). Let \(Q\) be a NAND program with \(n\) inputs, one
output, and \(m\) lines. We can assume without loss of generality that \(Q\)
contains the variables `one`

and `zero`

as usual.

We map \(Q\) to a \(3NAND\) formula \(\varphi\) as follows:

- \(\varphi\) has \(m+n\) variables \(z_0,\ldots,z_{m+n-1}\)
- For every \(\ell\in \{n,n+1,\ldots,n+m \}\), if the \(\ell-n\)-th line
of the program \(Q\) is
`foo := bar NAND blah`

then we add to \(\varphi\) the constraint \(z_\ell = NAND(z_j,z_k)\) where \(j-n\) and \(k-n\) correspond to the last lines in which the variables`bar`

and`blah`

(respectively) were written to. If one or both of`bar`

and`blah`

was not written to before then we use \(z_{\ell_0}\) instead of the corresponding value \(z_j\) or \(z_k\) in the constraint, where \(\ell_0-n\) is the line in which`zero`

is assigned a value. If one or both of`bar`

and`blah`

is an input variable`x_i`

then we we use \(z_i\) in the constraint. - Let \(\ell^*\) be the last line in which the output
`y_0`

is assigned a value. Then we add the constraint \(z_{\ell^*} = NAND(z_{\ell_0},z_{\ell_0})\) where \(\ell_0-n\) is as above the last line in which`zero`

is assigned a value. Note that this is effectively the constraint \(z_{\ell^*}=NAND(0,0)=1\).

To complete the proof we need to show that there exists \(w\in \{0,1\}^n\) s.t. \(Q(w)=1\) if and only if there exists \(z\in \{0,1\}^{n+m}\) that satisfies all constraints in \(\varphi\). We now show both sides of this equivalence.

**Completeness:**Suppose that there is \(w\in \{0,1\}^n\) s.t. \(Q(w)=1\). Let \(z\in \{0,1\}^{n+m}\) be defined as follows: for \(i\in [n]\), \(z_i=w_i\) and for \(i\in \{n,n+1,\ldots,n+m\}\) \(z_i\) equals the value that is assigned in the \((i-n)\)-th line of \(Q\) when executed on \(w\). Then by construction \(z\) satisfies all of the constraints of \(\varphi\) (including the constraint that \(z_{\ell^*}=NAND(0,0)=1\) since \(Q(w)=1\).)**Soundness:**Suppose that there exists \(z\in \{0,1\}^{n+m}\) satisfying \(\varphi\). Soundness will follow by showing that \(Q(z_0,\ldots,z_{n-1})=1\) (and hence in particular there exists \(w\in \{0,1\}^n\), namely \(w=z_0\cdots z_{n-1}\), such that \(Q(w)=1\)). To do this we will prove the following claim \((*)\): for every \(\ell \in [m]\), \(z_{\ell+n}\) equals the value assigned in the \(\ell\)-th step of the execution of the program \(Q\) on \(z_0,\ldots,z_{n-1}\). Note that because \(z\) satisfies the constraints of \(\varphi\), \((*)\) is sufficient to prove the soundness condition since these constraints imply that the last value assigned to the variable`y_0`

in the execution of \(Q\) on \(z_0\cdots w_{n-1}\) is equal to \(1\). To prove \((*)\) suppose, towards a contradiction, that it is false, and let \(\ell\) be the smallest number such that \(z_{\ell+n}\) is*not*equal to the value assigned in the \(\ell\)-th step of the exeuction of \(Q\) on \(z_0,\ldots,z_{n-1}\). But since \(z\) satisfies the constraints of \(\varphi\), we get that \(z_{\ell+n}=NAND(z_i,z_j)\) where (by the assumption above that \(\ell\) is*smallest*with this property) these values*do*correspond to the values last assigned to the variables on the righthand side of the assignment operator in the \(\ell\)-th line of the program. But this means that the value assigned in the \(\ell\)-th step is indeed simply the NAND of \(z_i\) and \(z_j\), contradicting our assumption on the choice of \(\ell\).

To conclude the proof of cook-levin-thm, we need to show threenand-sat-thm and show that \(3NAND \leq_p 3SAT\). We now do so.

To prove threenand-sat-thm we need to map a 3NAND formula \(\varphi\) into a 3SAT formula \(\psi\) such that \(\varphi\) is satisfiable if and only if \(\psi\) is. The idea is that we can transform every NAND constraint of the form \(a=NAND(b,c)\) into the AND of ORs involving the variables \(a,b,c\) and their negations, where each of the ORs contains at most three terms. The construction is fairly straightforward, and the details are given below.

It is a good exercise for you to try to find a 3CNF formula \(\xi\) on three variables \(a,b,c\) such that \(\xi(a,b,c)\) is true if and only if \(a = NAND(b,c)\). Once you do so, try to see why this implies a reduction from \(3NAND\) to \(3SAT\), and hence completes the proof of threenand-sat-thm

The constraint \[
z_i = NAND(z_j,z_k) \label{eq:NANDconstraint}
\] is satisfied if \(z_i=1\) whenever \((z_j,z_k) \neq (1,1)\). By going
through all cases, we can verify that \eqref{eq:NANDconstraint} is
equivalent to the constraint \[
(\overline{z_i} \vee \overline{z_j} \vee\overline{z_k} ) \wedge (z_i \vee z_j )
\wedge (z_i \vee z_k) \;\;. \label{eq:CNFNAND}
\] Indeed if \(z_j=z_k=1\) then the first constraint of
eq:CNFNAND is only true if \(z_i=0\). On the other hand, if
either of \(z_j\) or \(z_k\) equals \(0\) then unless \(z_i=1\) either the
second or third constraints will fail. This means that, given any 3NAND
formula \(\varphi\) over \(n\) variables \(z_0,\ldots,z_{n-1}\), we can obtain
a 3SAT formula \(\psi\) over the same variables by replacing every \(3NAND\)
constraint of \(\varphi\) with three \(3OR\) constraints as in
eq:CNFNAND.

We have shown that for every function \(F\) in \(\mathbf{NP}\), \(F \leq_p NANDSAT \leq_p 3NAND \leq_p 3SAT\), and so \(3SAT\) is \(\mathbf{NP}\)-hard. Since in reductionchap we saw that \(3SAT \leq_p QUADEQ\), \(3SAT \leq_p ISET\), \(3SAT \leq_p MAXCUT\) and \(3SAT \leq_p LONGPATH\), all these problems are \(\mathbf{NP}\)-hard as well. Finally, since all the aforementioned problems are in \(\mathbf{NP}\), they are all in fact \(\mathbf{NP}\)-complete and have equivalent complexity. There are thousands of other natural problems that are \(\mathbf{NP}\)-complete as well. Finding a polynomial-time algorithm for any one of them will imply a polynomial-time algorithm for all of them.

- Many of the problems for which we don’t know polynomial-time
algorithms are \(\mathbf{NP}\)-complete, which means that finding a
polynomial-time algorithm for one of them would imply a
polynomial-time algorithm for
*all*of them. - It is conjectured that \(\mathbf{NP}\neq \mathbf{P}\) which means that
we believe that polynomial-time algorithms for these problems are
not merely
*unknown*but are*nonexistent*. - While an \(\mathbf{NP}\)-hardness result means for example that a
full-fledged “textbook” solution to a problem such as MAX-CUT that
is as clean and general as the algorithm for MIN-CUT probably does
not exist, it does not mean that we need to give up whenever we see
a MAX-CUT instance. Later in this course we will discuss several
strategies to deal with \(\mathbf{NP}\)-hardness, including
*average-case complexity*and*approximation algorithms*.

Most of the exercises have been written in the summer of 2018 and haven’t yet been fully debugged. While I would prefer people do not post online solutions to the exercises, I would greatly appreciate if you let me know of any bugs. You can do so by posting a GitHub issue about the exercise, and optionally complement this with an email to me with more details about the attempted solution.

Prove that if there is no \(n^{O(\log^2 n)}\) time algorithm for \(3SAT\)
then there is some \(F\in \mathbf{NP}\) such that \(F \not\in \mathbf{P}\)
and \(F\) is not \(\mathbf{NP}\) complete.**Hint:** Use the function \(F\) that on input a formula \(\varphi\)
and a string of the form \(1^t\), outputs \(1\) if and only if \(\varphi\)
is satisfiable and \(t=|\varphi|^{\log|\varphi|}\).

Some topics related to this chapter that might be accessible to advanced students include: (to be completed)

Copyright 2018, Boaz Barak.

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