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Unformatted text preview: Chapter 1
Introduction: Some
Representative Problems
1.1 A First Problem: Stable Matching
As an opening topic, we look at an algorithmic problem that nicely illustrates
many of the themes we will be emphasizing. It is motivated by some very
natural and practical concerns, and from these we formulate a clean and
simple statement of a problem. The algorithm to solve the problem is very
clean as well, and most of our work will be spent in proving that it is correct
and giving an acceptable bound on the amount of time it takes to terminate
with an answer. The problem itself—the Stable Matching Problem—has several
origins. The Problem
The Stable Matching Problem originated, in part, in 1962, when David Gale
and Lloyd Shapley, two mathematical economists, asked the question: Could
one design a college admissions process, or a job recruiting process, that was
selfenforcing ? What did they mean by this?
To set up the question, let’s ﬁrst think informally about the kind of situation
that might arise as a group of friends, all juniors in college majoring in
computer science, begin applying to companies for summer internships. The
crux of the application process is the interplay between two different types
of parties: companies (the employers) and students (the applicants). Each
applicant has a preference ordering on companies, and each company—once
the applications come in—forms a preference ordering on its applicants. Based
on these preferences, companies extend offers to some of their applicants,
applicants choose which of their offers to accept, and people begin heading
off to their summer internships. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 1 (chap01) Windfall Software, PCA ZzTEX 11.6 2 Chapter 1 Introduction: Some Representative Problems Gale and Shapley considered the sorts of things that could start going
wrong with this process, in the absence of any mechanism to enforce the status
quo. Suppose, for example, that your friend Raj has just accepted a summer job
at the large telecommunications company CluNet. A few days later, the small
startup company WebExodus, which had been dragging its feet on making a
few ﬁnal decisions, calls up Raj and offers him a summer job as well. Now, Raj
actually prefers WebExodus to CluNet—won over perhaps by the laidback,
anythingcanhappen atmosphere—and so this new development may well
cause him to retract his acceptance of the CluNet offer and go to WebExodus
instead. Suddenly down one summer intern, CluNet offers a job to one of its
waitlisted applicants, who promptly retracts his previous acceptance of an
offer from the software giant Babelsoft, and the situation begins to spiral out
of control.
Things look just as bad, if not worse, from the other direction. Suppose
that Raj’s friend Chelsea, destined to go to Babelsoft but having just heard Raj’s
story, calls up the people at WebExodus and says, “You know, I’d really rather
spend the summer with you guys than at Babelsoft.” They ﬁnd this very easy
to believe; and furthermore, on looking at Chelsea’s application, they realize
that they would have rather hired her than some other student who actually
is scheduled to spend the summer at WebExodus. In this case, if WebExodus
were a slightly less scrupulous company, it might well ﬁnd some way to retract
its offer to this other student and hire Chelsea instead.
Situations like this can rapidly generate a lot of chaos, and many people—
both applicants and employers—can end up unhappy with the process as well
as the outcome. What has gone wrong? One basic problem is that the process
is not selfenforcing—if people are allowed to act in their selfinterest, then it
risks breaking down.
We might well prefer the following, more stable situation, in which selfinterest itself prevents offers from being retracted and redirected. Consider
another student, who has arranged to spend the summer at CluNet but calls
up WebExodus and reveals that he, too, would rather work for them. But in
this case, based on the offers already accepted, they are able to reply, “No, it
turns out that we prefer each of the students we’ve accepted to you, so we’re
afraid there’s nothing we can do.” Or consider an employer, earnestly following
up with its top applicants who went elsewhere, being told by each of them,
“No, I’m happy where I am.” In such a case, all the outcomes are stable—there
are no further outside deals that can be made.
So this is the question Gale and Shapley asked: Given a set of preferences
among employers and applicants, can we assign applicants to employers so
that for every employer E , and every applicant A who is not scheduled to work
for E , one of the following two things is the case? Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 2 (chap01) Windfall Software, PCA ZzTEX 11.6 3 1.1 A First Problem: Stable Matching (i) E prefers every one of its accepted applicants to A; or
(ii) A prefers her current situation over working for employer E .
If this holds, the outcome is stable: individual selfinterest will prevent any
applicant/employer deal from being made behind the scenes.
Gale and Shapley proceeded to develop a striking algorithmic solution to
this problem, which we will discuss presently. Before doing this, let’s note that
this is not the only origin of the Stable Matching Problem. It turns out that for
a decade before the work of Gale and Shapley, unbeknownst to them, the
National Resident Matching Program had been using a very similar procedure,
with the same underlying motivation, to match residents to hospitals. Indeed,
this system, with relatively little change, is still in use today.
This is one testament to the problem’s fundamental appeal. And from the
point of view of this book, it provides us with a nice ﬁrst domain in which
to reason about some basic combinatorial deﬁnitions and the algorithms that
build on them.
Formulating the Problem To get at the essence of this concept, it helps to
make the problem as clean as possible. The world of companies and applicants
contains some distracting asymmetries. Each applicant is looking for a single
company, but each company is looking for many applicants; moreover, there
may be more (or, as is sometimes the case, fewer) applicants than there are
available slots for summer jobs. Finally, each applicant does not typically apply
to every company.
It is useful, at least initially, to eliminate these complications and arrive at a
more “barebones” version of the problem: each of n applicants applies to each
of n companies, and each company wants to accept a single applicant. We will
see that doing this preserves the fundamental issues inherent in the problem;
in particular, our solution to this simpliﬁed version will extend directly to the
more general case as well.
Following Gale and Shapley, we observe that this special case can be
viewed as the problem of devising a system by which each of n men and
n women can end up getting married: our problem naturally has the analogue
of two “genders”—the applicants and the companies—and in the case we are
considering, everyone is seeking to be paired with exactly one individual of
the opposite gender.1 1 Gale and Shapley considered the samesex stable matching problem as well, where there is only a single gender. This is motivated by related applications, but it turns out to be fairly diﬀerent at a
technical level. Given the applicantemployer application we’re considering here, we’ll be focusing
on the version with two genders. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 3 (chap01) Windfall Software, PCA ZzTEX 11.6 4 Chapter 1 Introduction: Some Representative Problems So consider a set M = {m1, . . . , mn } of n men, and a set W = {w1, . . . , wn }
of n women. Let M × W denote the set of all possible ordered pairs of the form
(m , w), where m ∈ M and w ∈ W . A matching S is a set of ordered pairs, each
from M × W , with the property that each member of M and each member of
W appears in at most one pair in S. A perfect matching S is a matching with
the property that each member of M and each member of W appears in exactly
one pair in S .
An instability: m and w
each prefer the other to
their current partners. m w m w Figure 1.1 Perfect matching
S with instability (m , w ). Matchings and perfect matchings are objects that will recur frequently
throughout the book; they arise naturally in modeling a wide range of algorithmic problems. In the present situation, a perfect matching corresponds
simply to a way of pairing off the men with the women, in such a way that
everyone ends up married to somebody, and nobody is married to more than
one person—there is neither singlehood nor polygamy.
Now we can add the notion of preferences to this setting. Each man m ∈ M
ranks all the women; we will say that m prefers w to w if m ranks w higher
than w . We will refer to the ordered ranking of m as his preference list . We will
not allow ties in the ranking. Each woman, analogously, ranks all the men.
Given a perfect matching S, what can go wrong? Guided by our initial
motivation in terms of employers and applicants, we should be worried about
the following situation: There are two pairs (m , w) and (m , w ) in S (as
depicted in Figure 1.1) with the property that m prefers w to w, and w prefers
m to m .
In this case, there’s nothing to stop m and w from abandoning their current
partners and heading off together; the set of marriages is not selfenforcing.
We’ll say that such a pair (m , w ) is an instability with respect to S: (m , w )
does not belong to S, but each of m and w prefers the other to their partner
in S.
Our goal, then, is a set of marriages with no instabilities. We’ll say that
a matching S is stable if (i) it is perfect, and (ii) there is no instability with
respect to S. Two questions spring immediately to mind:
. Does there exist a stable matching for every set of preference lists? . Given a set of preference lists, can we efﬁciently construct a stable
matching if there is one? Some Examples To illustrate these deﬁnitions, consider the following two
very simple instances of the Stable Matching Problem.
First, suppose we have a set of two men, {m , m }, and a set of two women,
{w , w }. The preference lists are as follows:
m prefers w to w .
m prefers w to w . Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 4 (chap01) Windfall Software, PCA ZzTEX 11.6 5 1.1 A First Problem: Stable Matching w prefers m to m .
w prefers m to m .
If we think about this set of preference lists intuitively, it represents complete
agreement: the men agree on the order of the women, and the women agree
on the order of the men. There is a unique stable matching here, consisting
of the pairs (m , w) and (m , w ). The other perfect matching, consisting of the
pairs (m , w) and (m , w ), would not be a stable matching, because the pair
(m , w) would form an instability with respect to this matching. (Both m and
w would want to leave their respective partners and pair up.)
Next, here’s an example where things are a bit more intricate. Suppose
the preferences are
m prefers w to w .
m prefers w to w.
w prefers m to m.
w prefers m to m .
What’s going on in this case? The two men’s preferences mesh perfectly with
each other (they rank different women ﬁrst), and the two women’s preferences
likewise mesh perfectly with each other. But the men’s preferences clash
completely with the women’s preferences.
In this second example, there are two different stable matchings. The
matching consisting of the pairs (m , w) and (m , w ) is stable, because both
men are as happy as possible, so neither would leave their matched partner.
But the matching consisting of the pairs (m , w) and (m , w ) is also stable, for
the complementary reason that both women are as happy as possible. This is
an important point to remember as we go forward—it’s possible for an instance
to have more than one stable matching. Designing the Algorithm
We now show that there exists a stable matching for every set of preference
lists among the men and women. Moreover, our means of showing this will
also answer the second question that we asked above: we will give an efﬁcient
algorithm that takes the preference lists and constructs a stable matching.
Let us consider some of the basic ideas that motivate the algorithm.
. Initially, everyone is unmarried. Suppose an unmarried man m chooses
the woman w who ranks highest on his preference list and proposes to
her. Can we declare immediately that (m , w) will be one of the pairs in our
ﬁnal stable matching? Not necessarily: at some point in the future, a man
m whom w prefers may propose to her. On the other hand, it would be Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 5 (chap01) Windfall Software, PCA ZzTEX 11.6 6 Chapter 1 Introduction: Some Representative Problems dangerous for w to reject m right away; she may never receive a proposal
from someone she ranks as highly as m. So a natural idea would be to
have the pair (m , w) enter an intermediate state—engagement . Woman w will become
engaged to m if she
prefers him to m .
. m w m Figure 1.2 An intermediate
state of the GS algorithm
when a free man m is proposing to a woman w. Suppose we are now at a state in which some men and women are free —
not engaged—and some are engaged. The next step could look like this.
An arbitrary free man m chooses the highestranked woman w to whom
he has not yet proposed, and he proposes to her. If w is also free, then m
and w become engaged. Otherwise, w is already engaged to some other
man m . In this case, she determines which of m or m ranks higher
on her preference list; this man becomes engaged to w and the other
becomes free. . Finally, the algorithm will terminate when no one is free; at this moment,
all engagements are declared ﬁnal, and the resulting perfect matching is
returned. Here is a concrete description of the GaleShapley algorithm, with Figure 1.2 depicting a state of the algorithm.
Initially all m ∈ M and w ∈ W are free
While there is a man m who is free and hasn’t proposed to
every woman
Choose such a man m
Let w be the highestranked woman in m’s preference list
to whom m has not yet proposed
If w is free then
(m , w) become engaged
Else w is currently engaged to m
If w prefers m to m then
m remains free
Else w prefers m to m
(m , w) become engaged
m becomes free
Endif
Endif
Endwhile
Return the set S of engaged pairs An intriguing thing is that, although the GS algorithm is quite simple
to state, it is not immediately obvious that it returns a stable matching, or
even a perfect matching. We proceed to prove this now, through a sequence
of intermediate facts. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 6 (chap01) Windfall Software, PCA ZzTEX 11.6 7 1.1 A First Problem: Stable Matching Analyzing the Algorithm
First consider the view of a woman w during the execution of the algorithm.
For a while, no one has proposed to her, and she is free. Then a man m may
propose to her, and she becomes engaged. As time goes on, she may receive
additional proposals, accepting those that increase the rank of her partner. So
we discover the following.
(1.1) w remains engaged from the point at which she receives her ﬁrst
proposal; and the sequence of partners to which she is engaged gets better and
better (in terms of her preference list).
The view of a man m during the execution of the algorithm is rather
different. He is free until he proposes to the highestranked woman on his
list; at this point he may or may not become engaged. As time goes on, he
may alternate between being free and being engaged; however, the following
property does hold.
(1.2) The sequence of women to whom m proposes gets worse and worse (in
terms of his preference list).
Now we show that the algorithm terminates, and give a bound on the
maximum number of iterations needed for termination.
(1.3)
loop. The GS algorithm terminates after at most n2 iterations of the While Proof. A useful strategy for upperbounding the running time of an algorithm,
as we are trying to do here, is to ﬁnd a measure of progress. Namely, we seek
some precise way of saying that each step taken by the algorithm brings it
closer to termination.
In the case of the present algorithm, each iteration consists of some man
proposing (for the only time) to a woman he has never proposed to before. So
if we let P(t ) denote the set of pairs (m , w) such that m has proposed to w by
the end of iteration t , we see that for all t , the size of P(t + 1) is strictly greater
than the size of P(t ). But there are only n2 possible pairs of men and women
in total, so the value of P(·) can increase at most n2 times over the course of
the algorithm. It follows that there can be at most n2 iterations.
Two points are worth noting about the previous fact and its proof. First,
there are executions of the algorithm (with certain preference lists) that can
involve close to n2 iterations, so this analysis is not far from the best possible.
Second, there are many quantities that would not have worked well as a
progress measure for the algorithm, since they need not strictly increase in each Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 7 (chap01) Windfall Software, PCA ZzTEX 11.6 8 Chapter 1 Introduction: Some Representative Problems iteration. For example, the number of free individuals could remain constant
from one iteration to the next, as could the number of engaged pairs. Thus,
these quantities could not be used directly in giving an upper bound on the
maximum possible number of iterations, in the style of the previous paragraph.
Let us now establish that the set S returned at the termination of the
algorithm is in fact a perfect matching. Why is this not immediately obvious?
Essentially, we have to show that no man can “fall off” the end of his preference
list; the only way for the While loop to exit is for there to be no free man. In
this case, the set of engaged couples would indeed be a perfect matching.
So the main thing we need to show is the following.
(1.4) If m is free at some point in the execution of the algorithm, then there
is a woman to whom he has not yet proposed.
Proof. Suppose there comes a point when m is free but has already proposed
to every woman. Then by (1.1), each of the n women is engaged at this point
in time. Since the set of engaged pairs forms a matching, there must also be
n engaged men at this point in time. But there are only n men total, and m is
not engaged, so this is a contradiction.
(1.5) The set S returned at termination is a perfect matching. Proof. The set of engaged pairs always forms a matching. Let us suppose that
the algorithm terminates with a free man m. At termination, it must be the
case that m had already proposed to every woman, for otherwise the While
loop would not have exited. But this contradicts (1.4), which says that there
cannot be a free man who has proposed to every woman.
Finally, we prove the main property of the algorithm—namely, that it
results in a stable matching.
(1.6) Consider an execution of the GS algorithm that returns a set of pairs
S. The set S is a stable matching.
Proof. We have already seen, in (1.5), that S is a perfect matching. Thus, to
prove S is a stable matching, we will assume that there is an instability with
respect to S and obtain a contradiction. As deﬁned above, such an instability
would involve two pairs, (m , w) and (m , w ), in S with the properties that
. m prefers w to w, and . w prefers m to m . In the execution of the algorithm that produced S, m’s last proposal was, by
deﬁnition, to w. Now we ask: Did m propose to w at some earlier point in Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 8 (chap01) Windfall Software, PCA ZzTEX 11.6 9 1.1 A First Problem: Stable Matching this execution? If he didn’t, then w must occur higher on m’s preference list
than w , contradicting our assumption that m prefers w to w. If he did, then
he was rejected by w in favor of some other man m , whom w prefers to m.
m is the ﬁnal partner of w , so either m = m or, by (1.1), w prefers her ﬁnal
partner m to m ; either way this contradicts our assumption that w prefers
m to m .
It follows that S is a stable matching. Extensions
We began by deﬁning the notion of a stable matching; we have just proven
that the GS algorithm actually constructs one. We now consider some further
questions about the behavior of the GS algorithm and its relation to the
properties of different stable matchings.
To begin with, recall that we saw an example earlier in which there could
be multiple stable matchings; to recap, the preference lists in this example
were as follows:
m prefers w to w .
m prefers w to w.
w prefers m to m.
w prefers m to m .
Now, in any execution of the GaleShapley algorithm, m will become engaged
to w, m will become engaged to w (perhaps in the other order), and things
will stop there. Thus, the other stable matching, consisting of the pairs (m , w)
and (m , w ), is not attainable from an execution of the GS algorithm in which
the men propose. On the other hand, it would be reached if we ran a version of
the algorithm in which the women propose. And in larger examples, with more
than two people on each side, we can have an even larger collection of possible
stable matchings, many of them not achievable by any natural algorithm.
This example shows a certain “unfairness” in the GS algorithm, favoring
men. If the men’s preferences mesh perfectly (they all list different women as
their ﬁrst choice), then in all runs of the GS algorithm all men end up matched
with their ﬁrst choice, independent of the preferences of women. If women’s
preferences clash completely with the men’s preferences (as was the case in
this example), then the resulting stable matching is as bad as possible for the
women. So this simple set of preference lists compactly summarizes a world
in which someone is destined to end up unhappy: women are unhappy if men
propose, and men are unhappy if women propose.
Let’s now analyze the GS algorithm in more detail and try to understand
how general this “unfairness” phenomenon is. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 9 (chap01) Windfall Software, PCA ZzTEX 11.6 10 Chapter 1 Introduction: Some Representative Problems To begin with, our example reinforces the point that the GS algorithm
is actually underspeciﬁed: as long as there is a free man, we are allowed to
choose any free man to make the next proposal. Different choices specify
different executions of the algorithm; this is why, to be careful, we stated (1.6)
as “Consider an execution of the GS algorithm that returns a set of pairs S,”
instead of “Consider the set S returned by the GS algorithm.”
Thus, we encounter another very natural question: Do all executions of
the GS algorithm yield the same matching? This is a genre of question that
arises in many settings in computer science: we have an algorithm that runs
asynchronously, with different independent components performing actions
that can be interleaved in complex ways, and we want to know how much
variability this asynchrony causes in the ﬁnal outcome. To consider a very
different kind of example, the independent components may not be men and
women but electronic components activating parts of an airplane wing; the
effect of asynchrony in their behavior can be a big deal.
In the present context, we will see that the answer to our question is
surprisingly clean: all executions of the GS algorithm yield the same matching.
We proceed to prove this now.
All Executions Yield the Same Matching There are a number of possible
ways to prove a statement such as this, many of which would result in quite
complicated arguments. It turns out that the easiest and most informative approach for us will be to uniquely characterize the matching that is obtained and
then show that all executions result in the matching with this characterization.
What is the characterization? We’ll show that each man ends up with the
“best possible partner” in a concrete sense. (Recall that this is true if all men
prefer different women.) First, we will say that a woman w is a valid partner
of a man m if there is a stable matching that contains the pair (m , w). We will
say that w is the best valid partner of m if w is a valid partner of m, and no
woman whom m ranks higher than w is a valid partner of his. We will use
best (m) to denote the best valid partner of m.
Now, let S∗ denote the set of pairs {(m , best (m)) : m ∈ M }. We will prove
the following fact.
(1.7) Every execution of the GS algorithm results in the set S∗. This statement is surprising at a number of levels. First of all, as deﬁned,
there is no reason to believe that S∗ is matching at all, let alone a stable
matching. After all, why couldn’t it happen that two men have the same best
valid partner? Secondly, the result shows that the GS algorithm gives the best
possible outcome for every man simultaneously; there is no stable matching
in which any of the men could have hoped to do better. And ﬁnally, it answers Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 10 (chap01) Windfall Software, PCA ZzTEX 11.6 1.1 A First Problem: Stable Matching 11 our question above by showing that the order of proposals in the GS algorithm
has absolutely no effect on the ﬁnal outcome.
Despite all this, the proof is not so difﬁcult.
Proof. Let us suppose, by way of contradiction, that some execution E of the
GS algorithm results in a matching S in which some man is paired with a
woman who is not his best valid partner. Since men propose in decreasing
order of preference, this means that some man is rejected by a valid partner
during the execution E of the algorithm. So consider the ﬁrst moment during
the execution E in which some man, say m, is rejected by a valid partner w.
Again, since men propose in decreasing order of preference, and since this is
the ﬁrst time such a rejection has occurred, it must be that w is m’s best valid
partner best (m).
The rejection of m by w may have happened either because m proposed
and was turned down in favor of w’s existing engagement, or because w broke
her engagement to m in favor of a better proposal. But either way, at this
moment w forms an engagement with a man m whom she prefers to m.
Since w is a valid partner of m, there exists a stable matching S containing
the pair (m , w). Now we ask: Who is m paired with in this matching? Suppose
it is a woman w = w.
Since the rejection of m by w was the ﬁrst rejection of a man by a valid
partner in the execution E , it must be that m had not been rejected by any valid
partner at the point in E when he became engaged to w. Since he proposed in
decreasing order of preference, and since w is clearly a valid partner of m , it
must be that m prefers w to w . But we have already seen that w prefers m to
m, for in execution E she rejected m in favor of m . Since (m , w) ∈ S , (m , w)
is an instability in S .
This contradicts our claim that S is stable and hence contradicts our initial
assumption.
So for the men, the GS algorithm is ideal. Unfortunately, the same cannot
be said for the women. For a woman w, we say that m is a valid partner if
there is a stable matching that contains the pair (m , w). We say that m is the
worst valid partner of w if m is a valid partner of w, and no man whom w
ranks lower than m is a valid partner of hers.
(1.8) In the stable matching S∗, each woman is paired with her worst valid
partner.
Proof. Suppose there was a pair (m , w) in S∗ so that m is not the worst valid
partner of w. Then there is a stable matching S in which w is paired with a
man m whom she likes less than m. In S , m is paired with a woman w = w; Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 11 (chap01) Windfall Software, PCA ZzTEX 11.6 12 Chapter 1 Introduction: Some Representative Problems since w is the best valid partner of m, and w is a valid partner of m, we see
that m prefers w to w .
But from this it follows that (m , w) is an instability in S , contradicting the
claim that S is stable and hence contradicting our initial assumption.
Thus, we ﬁnd that our simple example above, in which the men’s preferences clashed with the women’s, hinted at a very general phenomenon: for
any input, the side that does the proposing in the GS algorithm ends up with
the best possible stable matching (from their perspective), while the side that
does not do the proposing correspondingly ends up with the worst possible
stable matching. 1.2 Five Representative Problems
The Stable Matching Problem provides us with a rich example of the process
of algorithm design. For many problems, this design process involves a few
signiﬁcant steps: formulating the problem with enough mathematical precision
that we can ask a concrete question and start thinking about algorithms to solve
it; developing an algorithm for the problem; proving the algorithm is correct;
and giving a bound on the running time so as to establish the algorithm’s
efﬁciency.
This highlevel strategy is carried out in practice with the help of a few
fundamental design techniques, which are very useful in assessing the inherent
complexity of a problem and in formulating an algorithm to solve it. As in any
area, becoming familiar with these design techniques is a gradual process; but
with experience one can start recognizing problems as belonging to identiﬁable
genres and appreciating how subtle changes in the statement of a problem can
have an enormous effect on its complexity.
To get this discussion started, then, it helps to pick out a few representative milestones that we’ll be encountering in our study of algorithms: cleanly
formulated problems, all resembling one another at a general level, but differing greatly in their difﬁculty and in the kinds of approaches that one brings
to bear on them. The ﬁrst three will be solvable efﬁciently by a sequence of
increasingly subtle algorithmic techniques; the fourth marks a major turning
point in our discussion, serving as an example of a problem believed to be unsolvable by any efﬁcient algorithms; and the ﬁfth hints at a class of problems
believed to be harder still.
The problems are selfcontained and are all motivated by computing
applications. To talk about some of them, though, it will help to use the
terminology of graphs. While graphs are a common topic in earlier computer
science courses, we’ll be introducing them in a fair amount of depth in Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 12 (chap01) Windfall Software, PCA ZzTEX 11.6 1.2 Five Representative Problems 13 Chapter 3; due to their enormous expressive power, we’ll also be using them
extensively throughout the book. For the discussion here, it’s enough to think
of a graph G as simply a way of encoding pairwise relationships among a set
of objects. Thus, G consists of a pair of sets (V , E )—a collection V of nodes;
and a collection E of edges, each of which “joins” two of the nodes. We thus
represent an edge e ∈ E as a twoelement subset of V : e = {u , v} for some
u , v ∈ V , where we call u and v the ends of e . We typically draw graphs as in
Figure 1.3, with each node as a small circle and each edge as a line segment
joining its two ends.
Let’s now turn to a discussion of the ﬁve representative problems. Interval Scheduling
Consider the following very simple scheduling problem. You have a resource—
it may be a lecture room, a supercomputer, or an electron microscope—and
many people request to use the resource for periods of time. A request takes
the form: Can I reserve the resource starting at time s, until time f ? We will
assume that the resource can be used by at most one person at a time. A
scheduler wants to accept a subset of these requests, rejecting all others, so
that the accepted requests do not overlap in time. The goal is to maximize the
number of requests accepted. Figure 1.3 Two graphs, each
on four nodes. More formally, there will be n requests labeled 1, . . . , n, with each request
i specifying a start time si and a ﬁnish time fi . Naturally, we have si < fi for all
i. Two requests i and j are compatible if the requested intervals do not overlap:
that is, either request i is for an earlier time interval than request j (fi ≤ sj ),
or request i is for a later time than request j (fj ≤ si ). We’ll say more generally
that a subset A of requests is compatible if all pairs of requests i , j ∈ A, i = j are
compatible. The goal is to select a compatible subset of requests of maximum
possible size.
We illustrate an instance of this Interval Scheduling Problem in Figure 1.4.
Note that there is a single compatible set of size four, and this is the largest
compatible set. Figure 1.4 An instance of the Interval Scheduling Problem. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 13 (chap01) Windfall Software, PCA ZzTEX 11.6 14 Chapter 1 Introduction: Some Representative Problems We will see shortly that this problem can be solved by a very natural
algorithm that orders the set of requests according to a certain heuristic and
then “greedily” processes them in one pass, selecting as large a compatible
subset as it can. This will be typical of a class of greedy algorithms that we
will consider for various problems—myopic rules that process the input one
piece at a time with no apparent lookahead. When a greedy algorithm can be
shown to ﬁnd an optimal solution for all instances of a problem, it’s often fairly
surprising. We typically learn something about the structure of the underlying
problem from the fact that such a simple approach can be optimal. Weighted Interval Scheduling
In the Interval Scheduling Problem, we sought to maximize the number of
requests that could be accommodated simultaneously. Now, suppose more
generally that each request interval i has an associated value , or weight ,
vi > 0; we could picture this as the amount of money we will make from
the ith individual if we schedule his or her request. Our goal will be to ﬁnd a
compatible subset of intervals of maximum total value.
The case in which vi = 1 for each i is simply the basic Interval Scheduling
Problem; but the appearance of arbitrary values changes the nature of the
maximization problem quite a bit. Consider, for example, that if v1 exceeds
the sum of all other vi , then the optimal solution must include interval 1
regardless of the conﬁguration of the full set of intervals. So any algorithm
for this problem must be very sensitive to the values, and yet degenerate to a
method for solving (unweighted) interval scheduling when all the values are
equal to 1.
There appears to be no simple greedy rule that walks through the intervals
one at a time, making the correct decision in the presence of arbitrary values.
Instead, we employ a technique, dynamic programming , that builds up the
optimal value over all possible solutions in a compact, tabular way that leads
to a very efﬁcient algorithm. Bipartite Matching
When we considered the Stable Matching Problem, we deﬁned a matching to
be a set of ordered pairs of men and women with the property that each man
and each woman belong to at most one of the ordered pairs. We then deﬁned
a perfect matching to be a matching in which every man and every woman
belong to some pair.
We can express these concepts more generally in terms of graphs, and in
order to do this it is useful to deﬁne the notion of a bipartite graph. We say that
a graph G = (V , E ) is bipartite if its node set V can be partitioned into sets X Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 14 (chap01) Windfall Software, PCA ZzTEX 11.6 15 1.2 Five Representative Problems and Y in such a way that every edge has one end in X and the other end in Y .
A bipartite graph is pictured in Figure 1.5; often, when we want to emphasize
a graph’s “bipartiteness,” we will draw it this way, with the nodes in X and
Y in two parallel columns. But notice, for example, that the two graphs in
Figure 1.3 are also bipartite.
Now, in the problem of ﬁnding a stable matching, matchings were built
from pairs of men and women. In the case of bipartite graphs, the edges are
pairs of nodes, so we say that a matching in a graph G = (V , E ) is a set of edges
M ⊆ E with the property that each node appears in at most one edge of M .
M is a perfect matching if every node appears in exactly one edge of M . x1 y1 x2 y2 x3 y3 x4 y4 x5 y5 Figure 1.5 A bipartite graph. To see that this does capture the same notion we encountered in the Stable
Matching Problem, consider a bipartite graph G with a set X of n men, a set Y
of n women, and an edge from every node in X to every node in Y . Then the
matchings and perfect matchings in G are precisely the matchings and perfect
matchings among the set of men and women.
In the Stable Matching Problem, we added preferences to this picture. Here,
we do not consider preferences; but the nature of the problem in arbitrary
bipartite graphs adds a different source of complexity: there is not necessarily
an edge from every x ∈ X to every y ∈ Y , so the set of possible matchings has
quite a complicated structure. In other words, it is as though only certain pairs
of men and women are willing to be paired off, and we want to ﬁgure out
how to pair off many people in a way that is consistent with this. Consider,
for example, the bipartite graph G in Figure 1.5: there are many matchings in
G , but there is only one perfect matching. (Do you see it?)
Matchings in bipartite graphs can model situations in which objects are
being assigned to other objects. Thus, the nodes in X can represent jobs, the
nodes in Y can represent machines, and an edge (xi , yj ) can indicate that
machine yj is capable of processing job xi . A perfect matching is then a way
of assigning each job to a machine that can process it, with the property that
each machine is assigned exactly one job. In the spring, computer science
departments across the country are often seen pondering a bipartite graph in
which X is the set of professors in the department, Y is the set of offered
courses, and an edge (xi .yj ) indicates that professor xi is capable of teaching
course yj . A perfect matching in this graph consists of an assignment of each
professor to a course that he or she can teach, in such a way that every course
is covered.
Thus the Bipartite Matching Problem is the following: Given an arbitrary
bipartite graph G , ﬁnd a matching of maximum size. If X  = Y  = n, then there
is a perfect matching if and only if the maximum matching has size n. We will
ﬁnd that the algorithmic techniques discussed earlier do not seem adequate Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 15 (chap01) Windfall Software, PCA ZzTEX 11.6 16 Chapter 1 Introduction: Some Representative Problems for providing an efﬁcient algorithm for this problem. There is, however, a very
elegant and efﬁcient algorithm to ﬁnd the maximum matching; it inductively
builds up larger and larger matchings, selectively backtracking along the way.
This process is called augmentation, and it forms the central component in a
large class of efﬁciently solvable problems called network ﬂow problems.
1 2 Independent Set
3 4 6 5 7 Figure 1.6 A graph whose
largest independent set has
size four. Now let’s talk about an extremely general problem, which includes most of
these earlier problems as special cases. Given a graph G = (V , E ), we say
a set of nodes S ⊆ V is independent if no two nodes in S are joined by an
edge. The Independent Set Problem is, then, the following: Given G , ﬁnd an
independent set that is as large as possible. For example, the maximum size of
an independent set in the graph in Figure 1.6 is four, achieved by the fournode
independent set {1, 4, 5, 6}.
The Independent Set Problem encodes any situation in which you are
trying to choose from among a collection of objects and there are pairwise
conﬂicts among some of the objects. Say you have n friends, and some pairs
of them don’t get along. How large a group of your friends can you invite to
dinner if you don’t want there to be any interpersonal tensions? This is simply
the largest independent set in the graph whose nodes are your friends, with
an edge between each conﬂicting pair.
Interval Scheduling and Bipartite Matching can both be encoded as special
cases of the Independent Set Problem. For Interval Scheduling, deﬁne a graph
G = (V , E ) in which the nodes are the intervals and there is an edge between
each pair of them that overlap; the independent sets in G are then just the
compatible subsets of intervals. Encoding Bipartite Matching as a special case
of Independent Set is a little trickier to see. Given a bipartite graph G = (V , E ),
the objects being chosen are edges, and the conﬂicts arise between two edges
that share an end. (These, indeed, are the pairs of edges that cannot belong
to a common matching.) So we deﬁne a graph G = (V , E ) in which the node
set V is equal to the edge set E of G . We deﬁne an edge between each pair
of elements in V that correspond to edges of G with a common end. We can
now check that the independent sets of G are precisely the matchings of G .
While it is not complicated to check this, it takes a little concentration to deal
with this type of “edgestonodes, nodestoedges” transformation.2 2 For those who are curious, we note that not every instance of the Independent Set Problem can arise in this way from Interval Scheduling or from Bipartite Matching; the full Independent Set Problem
really is more general. The ﬁrst graph in Figure 1.3 cannot arise as the “conﬂict graph” in an instance Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 16 (chap01) Windfall Software, PCA ZzTEX 11.6 1.2 Five Representative Problems 17 Given the generality of the Independent Set Problem, an efﬁcient algorithm
to solve it would be quite impressive. It would have to implicitly contain
algorithms for Interval Scheduling, Bipartite Matching, and a host of other
natural optimization problems.
The current status of Independent Set is this: no efﬁcient algorithm is
known for the problem, and it is conjectured that no such algorithm exists.
The obvious bruteforce algorithm would try all subsets of the nodes, checking
each to see if it is independent, and then recording the largest one encountered.
It is possible that this is close to the best we can do on this problem. We will
see later in the book that Independent Set is one of a large class of problems
that are termed NPcomplete . No efﬁcient algorithm is known for any of them;
but they are all equivalent in the sense that a solution to any one of them
would imply, in a precise sense, a solution to all of them.
Here’s a natural question: Is there anything good we can say about the
complexity of the Independent Set Problem? One positive thing is the following:
If we have a graph G on 1,000 nodes, and we want to convince you that it
contains an independent set S of size 100, then it’s quite easy. We simply
show you the graph G , circle the nodes of S in red, and let you check that
no two of them are joined by an edge. So there really seems to be a great
difference in difﬁculty between checking that something is a large independent
set and actually ﬁnding a large independent set. This may look like a very basic
observation—and it is—but it turns out to be crucial in understanding this class
of problems. Furthermore, as we’ll see next, it’s possible for a problem to be
so hard that there isn’t even an easy way to “check” solutions in this sense. Competitive Facility Location
Finally, we come to our ﬁfth problem, which is based on the following
twoplayer game. Consider two large companies that operate caf´ franchises
e
across the country—let’s call them JavaPlanet and Queequeg’s Coffee—and
they are currently competing for market share in a geographic area. First
JavaPlanet opens a franchise; then Queequeg’s Coffee opens a franchise; then
JavaPlanet; then Queequeg’s; and so on. . . . Suppose they must deal with
zoning regulations that require no two franchises be located too close together,
and each is trying to make its locations as convenient as possible. Who will
win?
Let’s make the rules of this “game” more concrete. The geographic region
in question is divided into n zones, labeled 1, 2, . . . , n. Each zone has a value of Interval Scheduling, and the second graph in Figure 1.3 cannot arise as the “conﬂict graph” in an
instance of Bipartite Matching. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 17 (chap01) Windfall Software, PCA ZzTEX 11.6 18 Chapter 1 10 Introduction: Some Representative Problems 1 5 15 5 1 5 1 15 10 Figure 1.7 An instance of the Competitive Facility Location Problem. bi , which is the revenue obtained by either of the companies if it opens a
franchise there. Finally, certain pairs of zones (i , j ) are adjacent , and local
zoning laws prevent two adjacent zones from each containing a franchise,
regardless of which company owns them. (They also prevent two franchises
from being opened in the same zone.) We model these conﬂicts via a graph
G = (V , E ), where V is the set of zones, and (i , j ) is an edge in E if the
zones i and j are adjacent. The zoning requirement then says that the full
set of franchises opened must form an independent set in G .
Thus our game consists of two players, P1 and P2, alternately selecting
nodes in G , with P1 moving ﬁrst. At all times, the set of all selected nodes
must form an independent set in G . Suppose that player P2 has a target bound
B, and we want to know: is there a strategy for P2 so that no matter how P1
plays, P2 will be able to select a set of nodes with a total value of at least B?
We will call this an instance of the Competitive Facility Location Problem.
Consider, for example, the instance pictured in Figure 1.7, and suppose
that P2’s target bound is B = 20. Then P2 does have a winning strategy. On the
other hand, if B = 25, then P2 does not.
One can work this out by looking at the ﬁgure for a while; but it requires
some amount of casechecking of the form, “If P1 goes here, then P2 will go
there; but if P1 goes over there, then P2 will go here. . . . ” And this appears to
be intrinsic to the problem: not only is it computationally difﬁcult to determine
whether P2 has a winning strategy; on a reasonablysized graph, it would even
be hard for us to convince you that P2 has a winning strategy. There does not
seem to be a short proof we could present; rather, we’d have to lead you on a
lengthy casebycase analysis of the set of possible moves.
This is in contrast to the Independent Set Problem, where we believe that
ﬁnding a large solution is hard but checking a proposed large solution is easy.
This contrast can be formalized in the class of PSPACEcomplete problems, of
which Competitive Facility Location is an example. PSPACEcomplete problems are believed to be strictly harder than NPcomplete problems, and this
conjectured lack of short “proofs” for their solutions is one indication of this
greater hardness. The notion of PSPACEcompleteness turns out to capture a
large collection of problems involving gameplaying and planning; many of
these are fundamental issues in the area of artiﬁcial intelligence. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 18 (chap01) Windfall Software, PCA ZzTEX 11.6 Solved Exercises 19 Solved Exercises
Solved Exercise 1
Consider a town with n men and n women seeking to get married to one
another. Each man has a preference list that ranks all the women, and each
woman has a preference list that ranks all the men.
The set of all 2n people is divided into two categories: good people and
bad people. Suppose that for some number k, 1 ≤ k ≤ n − 1, there are k good
men and k good women; thus there are n − k bad men and n − k bad women.
Everyone would rather marry any good person than any bad person.
Formally, each preference list has the property that it ranks each good person
of the opposite gender higher than each bad person of the opposite gender: its
ﬁrst k entries are the good people (of the opposite gender) in some order, and
its next n − k are the bad people (of the opposite gender) in some order.
Show that in every stable matching, every good man is married to a good
woman.
Solution A natural way to get started thinking about this problem is to
assume the claim is false and try to work toward obtaining a contradiction.
What would it mean for the claim to be false? There would exist some stable
matching M in which a good man m was married to a bad woman w.
Now, let’s consider what the other pairs in M look like. There are k good
men and k good women. Could it be the case that every good woman is married
to a good man in this matching M ? No: one of the good men (namely m) is
already married to a bad woman, and that leaves only k − 1 other good men.
So even if all of them were married to good women, that would still leave some
good woman who is married to a bad man.
Let w be such a good woman, and let m be the bad man she is married to.
It is now easy to identify an instability in M : consider the pair (m , w ). Each is
good, but is married to a bad partner. Thus, each of m and w prefers the other
to their current partner, and hence (m , w ) is an instability. This contradicts
our assumption that M is stable, and hence concludes the proof. Solved Exercise 2
We can think about a generalization of the Stable Matching Problem, in which
certain manwoman pairs are explicitly forbidden. In the case of employers and
applicants, we could imagine that certain applicants simply lack the necessary
qualiﬁcations or certiﬁcations; and so they cannot be employed at certain
companies, however desirable they may seem. Using the analogy to marriage
between men and women, we have a set M of n men, a set W of n women, Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 19 (chap01) Windfall Software, PCA ZzTEX 11.6 20 Chapter 1 Introduction: Some Representative Problems and a set F ⊆ M × W of pairs who are simply not allowed to get married. Each
man m ranks all the women w for which (m , w) ∈ F , and each woman w ranks
all the men m for which (m , w ) ∈ F .
In this more general setting, we say that a matching S is stable if it does
not exhibit any of the following types of instability.
(i) There are two pairs (m , w) and (m , w ) in S with the property that (m , w )
not in F , m prefers w to w, and w prefers m to m . (The usual kind of
instability.)
(ii) There is a pair (m , w) ∈ S, and a man m , so that m is not part of any
pair in the matching, (m , w) ∈ F , and w prefers m to m. (A single man
is more desirable and not forbidden.)
(iii) There is a pair (m , w) ∈ S, and a woman w , so that w is not part of
any pair in the matching, (m , w ) ∈ F , and m prefers w to w. (A single
woman is more desirable and not forbidden.)
(iv) There is a man m and a woman w, neither of whom is part of any pair
in the matching, so that (m , w) ∈ F . (There are two single people with
nothing preventing them from getting married to each other.)
Note that under these more general deﬁnitions, a stable matching need not be
a perfect matching.
Now we can ask: For every set of preference lists and every set of forbidden
pairs, is there always a stable matching? Resolve this question by doing one of
the following two things: (a) give an algorithm that, for any set of preference
lists and forbidden pairs, produces a stable matching; or (b) give an example
of a set of preference lists and forbidden pairs for which there is no stable
matching.
Solution The GaleShapley algorithm is remarkably robust to variations on
the Stable Matching problem. So, if you’re faced with a new variation of the
problem and can’t ﬁnd a counterexample to stability, it’s often a good idea to
check whether a direct adaptation of the GS algorithm will in fact produce
stable matchings.
That turns out to be the case here. We will show that there is always a
stable matching, even in this more general model with forbidden pairs, and
we will do this by adapting the GS algorithm. To do this, let’s consider why
the original GS algorithm can’t be used directly. The difﬁculty, of course, is
that the GS algorithm doesn’t know anything about forbidden pairs, and so
the condition in the While loop, While there is a man m who is free and hasn’t proposed to
every woman, Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 20 (chap01) Windfall Software, PCA ZzTEX 11.6
Solved Exercises 21 won’t work: we don’t want m to propose to a woman w for which the pair
(m , w) is forbidden.
Thus, let’s consider a variation of the GS algorithm in which we make
only one change: we modify the While loop to say, While there is a man m who is free and hasn’t proposed to
every woman w for which (m , w) ∈ F.
Here is the algorithm in full:
Initially all m ∈ M and w ∈ W are free
While there is a man m who is free and hasn’t proposed to
every woman w for which (m , w) ∈ F
Choose such a man m
Let w be the highestranked woman in m’s preference list
to which m has not yet proposed
If w is free then
(m , w) become engaged
Else w is currently engaged to m
If w prefers m to m then
m remains free
Else w prefers m to m
(m , w) become engaged
m becomes free
Endif
Endif
Endwhile
Return the set S of engaged pairs We now prove that this yields a stable matching, under our new deﬁnition
of stability.
To begin with, facts (1.1), (1.2), and (1.3) from the text remain true (in
particular, the algorithm will terminate in at most n2 iterations). Also, we
don’t have to worry about establishing that the resulting matching S is perfect
(indeed, it may not be). We also notice an additional pairs of facts. If m is a
man who is not part of a pair in S, then m must have proposed to every nonforbidden woman; and if w is a woman who is not part of a pair in S, then it
must be that no man ever proposed to w.
Finally, we need only show
(1.9) There is no instability with respect to the returned matching S. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 21 (chap01) Windfall Software, PCA ZzTEX 11.6 22 Chapter 1 Introduction: Some Representative Problems Proof. Our general deﬁnition of instability has four parts: This means that we
have to make sure that none of the four bad things happens.
First, suppose there is an instability of type (i), consisting of pairs (m , w)
and (m , w ) in S with the property that (m , w) not in F , and m prefers w to
w, and w prefers m to m . It follows that m must have proposed to w ; so w
rejected m, and thus she prefers her ﬁnal partner to m—a contradiction.
Next, suppose there is an instability of type (ii), consisting of a pair
(m , w) ∈ S, and a man m , so that m is not part of any pair in the matching,
(m , w) ∈ F , and w prefers m to m. Then m must have proposed to w and
been rejected; again, it follows that w prefers her ﬁnal partner to m —a
contradiction.
Third, suppose there is an instability of type (ii ), consisting of a pair
(m , w) ∈ S, and a woman w , so that w is not part of any pair in the matching,
(m , w ) ∈ F , and m prefers w to w. Then no man proposed to w at all;
in particular, m never proposed to w , and so he must prefer w to w —a
contradiction.
Finally, suppose there is an instability of type (iii), consisting of a man
m and a woman w, neither of which is part of any pair in the matching, so
that (m , w) ∈ F . But for m to be single, he must have proposed to every nonforbidden woman; in particular, he must have proposed to w, which means
she would no longer be single—a contradiction. Exercises
1. Decide whether you think the following statement is true or false. If it is
true, give a short explanation. If it is false, give a counterexample.
True or false? In every instance of the Stable Matching Problem, there is a
stable matching containing a pair (m , w) such that m is ranked ﬁrst on the
preference list of w and w is ranked ﬁrst on the preference list of m.
2. Decide whether you think the following statement is true or false. If it is
true, give a short explanation. If it is false, give a counterexample.
True or false? Consider an instance of the Stable Matching Problem in which
there exists a man m and a woman w such that m is ranked ﬁrst on the
preference list of w and w is ranked ﬁrst on the preference list of m. Then in
every stable matching S for this instance, the pair (m , w) belongs to S.
3. There are many other settings in which we can ask questions related
to some type of “stability” principle. Here’s one, involving competition
between two enterprises. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 22 (chap01) Windfall Software, PCA ZzTEX 11.6 Exercises 23 Suppose we have two television networks, whom we’ll call A and D.
There are n primetime programming slots, and each network has n TV
shows. Each network wants to devise a schedule —an assignment of each
show to a distinct slot—so as to attract as much market share as possible.
Here is the way we determine how well the two networks perform
relative to each other, given their schedules. Each show has a fixed rating ,
which is based on the number of people who watched it last year; we’ll
assume that no two shows have exactly the same rating. A network wins a
given time slot if the show that it schedules for the time slot has a larger
rating than the show the other network schedules for that time slot. The
goal of each network is to win as many time slots as possible.
Suppose in the opening week of the fall season, Network A reveals a
schedule S and Network D reveals a schedule T . On the basis of this pair
of schedules, each network wins certain time slots, according to the rule
above. We’ll say that the pair of schedules (S , T ) is stable if neither network
can unilaterally change its own schedule and win more time slots. That
is, there is no schedule S so that Network A wins more slots with the
pair (S , T ) than it did with the pair (S , T ); and symmetrically, there is no
schedule T so that Network D wins more slots with the pair (S , T ) than
it did with the pair (S , T ).
The analogue of Gale and Shapley’s question for this kind of stability
is: For every set of TV shows and ratings, is there always a stable pair
of schedules? Resolve this question by doing one of the following two
things: (a) give an algorithm that, for any set of TV shows and associated
ratings, produces a stable pair of schedules; or (b) give an example of a
set of TV shows and associated ratings for which there is no stable pair
of schedules.
4. Gale and Shapley published their paper on the stable marriage problem
in 1962; but a version of their algorithm had already been in use for
ten years by the National Resident Matching Program, for the problem of
assigning medical residents to hospitals.
Basically, the situation was the following. There were m hospitals,
each with a certain number of available positions for hiring residents.
There were n medical students graduating in a given year, each interested
in joining one of the hospitals. Each hospital had a ranking of the students
in order of preference, and each student had a ranking of the hospitals
in order of preference. We will assume that there were more students
graduating than there were slots available in the m hospitals.
The interest, naturally, was in finding a way of assigning each student
to at most one hospital, in such a way that all available positions in all Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 23 (chap01) Windfall Software, PCA ZzTEX 11.6 24 Chapter 1 Introduction: Some Representative Problems hospitals were filled. (Since we are assuming a surplus of students, there
would be some students who do not get assigned to any hospital.)
We say that an assignment of students to hospitals is stable if neither
of the following situations arises.
. . First type of instability: There are students s and s , and a hospital h,
so that
– s is assigned to h, and
– s is assigned to no hospital, and
– h prefers s to s.
Second type of instability: There are students s and s , and hospitals
h and h , so that
– s is assigned to h, and
– s is assigned to h , and
– h prefers s to s, and
– s prefers h to h . So we basically have the stable marriage problem from the section,
except that (i) hospitals generally want more than one resident, and (ii)
there is a surplus of medical students.
Show that there is always a stable assignment of students to hospitals, and give an efficient algorithm to find one.
5. The Stable Matching Problem, as discussed in the text, assumes that all
men and women have a fully ordered list of preferences. In this problem
we will consider a version of the problem in which men and women can be
indifferent between certain options. As before we have a set M of n men
and a set W of n women. Assume each man and each woman ranks the
members of the opposite gender, but now we allow ties in the ranking.
For example (with n = 4), a woman could say that m1 is ranked in first
place; second place is a tie between m2 and m3 (she has no preference
between them); and m4 is in last place. We will say that w prefers m to m
if m is ranked higher than m on her preference list (they are not tied).
With indifferences in the rankings, there could be two natural notions
for stability. And for each, we can ask about the existence of stable
matchings, as follows.
(a) Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 A strong instability in a perfect matching S consists of a man m and a
woman w, such that m and w prefer each other to their partners in S.
Does there always exist a perfect matching with no strong instability?
Either give an example of a set of men and women with preference
lists for which every perfect matching has a strong instability; or give p. 24 (chap01) Windfall Software, PCA ZzTEX 11.6 Exercises 25 an efficient algorithm that is guaranteed to find a perfect matching
with no strong instability.
(b) A weak instability in a perfect matching S is a man m and a woman
w, such that their partners in S are w and m , respectively, and one
of the following holds:
– m prefers w to w , and w either prefers m to m or is indifferent
between these two choices; or
– w prefers m to m , and m either prefers w to w or is indifferent
between these two choices.
In other words, the pairing between m and w is either preferred by
both, or preferred by one while the other is indifferent. Does there
always exists a perfect matching with no weak instability? Either
give an example of a set of men and women with preference lists
for which every perfect matching has a weak instability; or give an
efficient algorithm that is guaranteed to find a perfect matching with
no weak instability.
6. Peripatetic Shipping Lines, Inc., is a shipping company that owns n ships
and provides service to n ports. Each of its ships has a schedule which
says, for each day of the month, which of the ports it’s currently visiting,
or whether it’s out at sea. (You can assume the “month” here has m days,
for some m > n.) Each ship visits each port for exactly one day during the
month. For safety reasons, PSL Inc. has the following strict requirement:
(†) No two ships can be in the same port on the same day.
The company wants to perform maintenance on all the ships this
month, via the following scheme. They want to truncate each ship’s
schedule: for each ship Si , there will be some day when it arrives in its
scheduled port and simply remains there for the rest of the month (for
maintenance). This means that Si will not visit the remaining ports on
its schedule (if any) that month, but this is okay. So the truncation of
Si ’s schedule will simply consist of its original schedule up to a certain
specified day on which it is in a port P ; the remainder of the truncated
schedule simply has it remain in port P .
Now the company’s question to you is the following: Given the schedule for each ship, find a truncation of each so that condition (†) continues
to hold: no two ships are ever in the same port on the same day.
Show that such a set of truncations can always be found, and give an
efficient algorithm to find them. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 25 (chap01) Windfall Software, PCA ZzTEX 11.6 26 Chapter 1 Introduction: Some Representative Problems Example: Suppose we have two ships and two ports, and the “month” has
four days. Suppose the first ship’s schedule is
port P1; at sea; port P2; at sea
and the second ship’s schedule is
at sea; port P1; at sea; port P2
Then the (only) way to choose truncations would be to have the first ship
remain in port P2 starting on day 3, and have the second ship remain in
port P1 starting on day 2.
7. Some of your friends are working for CluNet, a builder of large communication networks, and they are looking at algorithms for switching in a
particular type of input/output crossbar.
Here is the setup. There are n input wires and n output wires , each
directed from a source to a terminus . Each input wire meets each output
wire in exactly one distinct point, at a special piece of hardware called
a junction box . Points on the wire are naturally ordered in the direction
from source to terminus; for two distinct points x and y on the same wire,
we say that x is upstream from y if x is closer to the source than y, and
otherwise we say x is downstream from y. The order in which one input
wire meets the output wires is not necessarily the same as the order in
which another input wire meets the output wires. (And similarly for the
orders in which output wires meet input wires.)
Now, here’s the switching component of this situation. Each input
wire is carrying a distinct data stream, and this data stream must be
switched onto one of the output wires. If the stream of Input i is switched
onto Output j , at junction box B, then this stream passes through all
junction boxes upstream from B on Input i, then through B, then through
all junction boxes downstream from B on Output j . It does not matter
which input data stream gets switched onto which output wire, but
each input data stream must be switched onto a different output wire.
Furthermore—and this is the tricky constraint—no two data streams can
pass through the same junction box following the switching operation.
Finally, here’s the problem. Show that for any specified pattern in
which the input wires and output wires meet each other (each pair meeting exactly once), a valid switching of the data streams can always be
found—one in which each input data stream is switched onto a different output, and no two of the resulting streams pass through the same
junction box. Additionally, give an efficient algorithm to find such a valid
switching. (Figure 1.8 gives an example with its solution.) Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 26 (chap01) Windfall Software, PCA ZzTEX 11.6 Exercises Junction
Junction Junction Junction 27 Output 1
(meets Input 2
before Input 1) Output 2
(meets Input 2
before Input 1) Input 1
Input 2
(meets Output 2 (meets Output 1
before Output 1) before Output 2)
Figure 1.8 An example with two input wires and two output wires. Input 1 has its
junction with Output 2 upstream from its junction with Output 1; Input 2 has its
junction with Output 1 upstream from its junction with Output 2. A valid solution is
to switch the data stream of Input 1 onto Output 2, and the data stream of Input 2
onto Output 1. On the other hand, if the stream of Input 1 were switched onto Output
1, and the stream of Input 2 were switched onto Output 2, then both streams would
pass through the junction box at the meeting of Input 1 and Output 2—and this is not
allowed. 8. For this problem, we will explore the issue of truthfulness in the Stable
Matching Problem and specifically in the GaleShapley algorithm. The
basic question is: Can a man or a woman end up better off by lying about
his or her preferences? More concretely, we suppose each participant has
a true preference order. Now consider a woman w. Suppose w prefers man
m to m , but both m and m are low on her list of preferences. Can it be the
case that by switching the order of m and m on her list of preferences (i.e.,
by falsely claiming that she prefers m to m) and running the algorithm
with this false preference list, w will end up with a man m that she truly
prefers to both m and m ? (We can ask the same question for men, but
will focus on the case of women for purposes of this question.)
Resolve this questions by doing one of the following two things:
(a) Give a proof that, for any set of preference lists, switching the
order of a pair on the list cannot improve a woman’s partner in the GaleShapley algorithm; or Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 27 (chap01) Windfall Software, PCA ZzTEX 11.6 28 Chapter 1 Introduction: Some Representative Problems (b) Give an example of a set of preference lists for which there is
a switch that would improve the partner of a woman who switched
preferences. Notes and Further Reading
The Stable Matching Problem was ﬁrst deﬁned and analyzed by Gale and
Shapley (1962); according to David Gale, their motivation for the problem
came from a story they had recently read in the New Yorker about the intricacies
of the college admissions process (Gale, 2001). Stable matching has grown
into an area of study in its own right, covered in books by Gusﬁeld and Irving
(1989) and Knuth (1997). Gusﬁeld and Irving also provide a nice survey of the
“parallel” history of the Stable Matching Problem as a technique invented for
matching applicants with employers in medicine and other professions.
As discussed in the chapter, our ﬁve representative problems will be
central to the book’s discussions, respectively, of greedy algorithms, dynamic
programming, network ﬂows, NPcompleteness, and PSPACEcompleteness.
We will discuss the problems in these contexts later in the book. Kleinberg & Tardos ﬁrst pages 2005/2/1 11:06 p. 28 (chap01) Windfall Software, PCA ZzTEX 11.6 ...
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This note was uploaded on 11/23/2009 for the course CS CMPT 307 taught by Professor A.bulatov during the Spring '09 term at Simon Fraser.
 Spring '09
 A.BULATOV
 Algorithms

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