Lecture26

Approximation Algorithms
Introduction

Why Approximation Algorithms
 Problems that we cannot find an optimal
solution in a polynomial time
 Eg: Set Cover, Bin Packing
 Need to find a near-optimal solution:
 Heuristic
 Approximation algorithms:
 This gives us a guarantee approximation ratio

Introduction to Combinatorial
Optimization

Combinatorial Optimization
 The study of finding the “best” object from
within some finite space of objects, eg:
 Shortest path: Given a graph with edge costs and a
pair of nodes, find the shortest path (least costs)
between them
 Traveling salesman: Given a complete graph with
nonnegative edge costs, find a minimum cost cycle
visiting every vertex exactly once
 Maximum Network Lifetime: Given a wireless
sensor networks and a set of targets, find a schedule
of these sensors to maximize network lifetime

In P or not in P?
Informal Definitions:
 The class P consists of those problems that are
solvable in polynomial time, i.e. O(nk) for some
constant k where n is the size of the input.
 The class NP consists of those problems that
are “verifiable” in polynomial time:
 Given a certificate of a solution, then we can verify
that the certificate is correct in polynomial time

In P or not in P: Examples
 In P:
 Shortest path
 Minimum Spanning Tree
 Not in P (NP):
 Vertex Cover
 Traveling salesman
 Minimum Connected Dominating Set

NP-completeness (NPC)
 A problem is in the class NPC if it is in NP and
is as “hard” as any problem in NP

What is “hard”
 Decision Problems: Only consider YES-NO
 Decision version of TSP: Given n cities, is there a
TSP tour of length at most L?
 Why Decision Problem? What relation between
the optimization problem and its decision?
 Decision is not “harder” than that of its optimization
problem
 If the decision is “hard”,then the optimization
problem should be “hard”

NP-complete and NP-hard
A language L is NP-complete if:
1. L is in NP, and
2. For every L` in NP, L` can be polynomial-time
reducible to L

Approximation Algorithms
 An algorithm that returns near-optimal
solutions in polynomial time
 Approximation Ratio ρ(n):
 Define: C* as a optimal solution and C is the
solution produced by the approximation algorithm
 max (C/C*, C*/C) <= ρ(n)
 Maximization problem: 0 < C <= C*, thus C*/C
shows that C* is larger than C by ρ(n)
 Minimization problem: 0 < C* <= C, thus C/C*
shows that C is larger than C* by ρ(n)

Approximation Algorithms (cont)
 PTAS (Polynomial Time Approximation
Scheme): A (1 + ε)-approximation algorithm
for a NP-hard optimization П where its running
time is bounded by a polynomial in the size of
instance I
 FPTAS (Fully PTAS): The same as above +
time is bounded by a polynomial in both the
size of instance I and 1/ε

A Dilemma!
 We cannot find C*, how can we compare C to
C*?
 How can we design an algorithm so that we can
compare C to C*
It is the objective of this course!!!

Techniques
 A variety of techniques to design and analyze
many approximation algorithms for
computationally hard problems:
 Combinatorial algorithms:
 Greedy Techniques, Independent System, Submodular
Function
 Linear Programming based algorithms
 Semidefinite Programming based algorithms

Vertex Cover
 Definition:
 An Example
'atincidentendpointone
leastathas,edgeeveryforiffofcoververtexa
calledis'subseta,),(graphundirectedanGiven
V
eEeG
VVEVG

Vertex Cover Problem
 Definition:
 Given an undirected graph G=(V,E), find a vertex
cover with minimum size (has the least number of
vertices)
 This is sometimes called cardinality vertex cover
 More generalization:
 Given an undirected graph G=(V,E), and a cost
function on vertices c: V → Q+
 Find a minimum cost vertex cover

How to solve it
 Matching:
 A set M of edges in a graph G is called a matching
of G if no two edges in set M have an endpoint in
common
 Example:

How to solve it (cont)
 Maximum Matching:
 A matching of G with the greatest number of edges
 Maximal Matching:
 A matching which is not contained in any larger
matching
 Note: Any maximum matching is certainly
maximal, but not the reverse

Main Observation
 No vertex can cover two edges of a matching
 The size of every vertex cover is at least the
size of every matching: |M| ≤ |C|
 |M| = |C| indicates the optimality
 Possible Solution: Using Maximal matching to
find Minimum vertex cover

An Algorithm
 Alg 1: Find a maximal matching in G and output the
set of matched vertices
 Theorem: Alg 1 is a factor 2-approximation algorithm.
 Proof:
optC
MCoptM
Copt
2||Therefore,
||2||and||
:haveWe1.algorithmfromobtainedcoververtex
ofsetabeLetsolution.optimalanofsizeabeLet

Can Alg1 be improved?
 Q1: Can the approximation ratio of Alg 1 be
improved by a better analysis?
 Q2: Can we design a better approximation
algorithm using the similar technique (maximal
matching)?
 Q3: Is there any other lower bounding method
that can lead to a better approximation
algorithm?

Answers
 A1: No by considering the complete bipartite
graphs Kn,n
 A2: No by considering the complete graph Kn
where n is odd.
 |M| = (n-1)/2 whereas opt = n -1

Answers (cont)
 A3:
 Currently a central open problem
 Yes, we can obtain a better one by using the
semidefinite programming
 Generalization vertex Cover
 Can we still able to design a 2-approximation
algorithm?
 Homework assignment!

Set Cover
 Definition: Given a universe U of n elements, a
collection of subsets of U, S = {S1, …, Sm}, and a cost
function c: S → Q+, find a minimum cost
subcollection C of S that covers all elements of U.
 Example:
 U = {1, 2, 3, 4, 5}
 S1 = {1, 2, 3}, S2 = {2,3}, S3 = {4, 5}, S4 = {1, 2, 4}
 c1 = c2 = c3 = c4 = 1
 Solution C = {S1, S3}
 If the cost is uniform, then the set cover problem asks
us to find a subcollection covering all elements of U
with the minimum size.

INSTANCE: Given a universe U of n elements, a collection
of subsets of U, S = {S1, …, Sm}, and a positive integer b
QUESTION: Is there a , |C| ≤ b,
such that
(Note: The subcollection {Si | } satisfying the above
condition is called a set cover of U
NP-completeness
 Theorem: Set Cover (SC) is NP-complete
 Proof:

Proof (cont)
 First we need to show that SC is in NP. Given a
subcollection C, it is easy to verify that if |C| ≤
b and the union of all sets listed in C does
include all elements in U.
 To complete the proof, we need to show that
Vertex Cover (VC) ≤p Set Cover (SC)
Given an instance C of VC (an undirected graph
G=(V,E) and a positive integer j), we need to
construct an instance C’ of SC in polynomial
time such that C is satisfiable iff C’ is
satisfiable.

Proof (cont)
Construction: Let U = E. We will define n
elements of U and a collection S as follows:
Note: Each edge corresponds to each element in U and
each vertex corresponds to each set in S.
Label all vertices in V from 1 to n. Let Si be the set of all
edges that incident to vertex i. Finally, let b = j. This
construction is in poly-time with respect to the size of
VC instance.

VERTEX-COVER p SET-COVER
one set for every vertex,
containing the edges it covers
VC
one element
for every edge
VC SC

Proof (cont)
Now, we need to prove that C is satisfiable iff C’ is.
That is, we need to show that if the original instance
of VC is a yes instance iff the constructed instance of
SC is a yes instance.
 (→) Suppose G has a vertex cover of size at most j,
called C. By our construction, C corresponds to a
collection C’ of subsets of U. Since b = j, |C’| ≤ b. Plus,
C’ covers all elements in U since C “covers” all edges
in G. To see this, consider any element of U. Such an
element is an edge in G. Since C is a set cover, at least
one endpoint of this edge is in C.

 (←) Suppose there is a set cover C’ of size at
most b in our constructed instance. Since each
set in C’ is associated with a vertex in G, let C
be the set of these vertices. Then |C| = |C’| ≤ b =
j. Plus, C is a vertex cover of G since C’ is a set
cover. To see this, consider any edge e. Since e
is in U, C’ must contain at least one set that
includes e. By our construction, the only set
that include e correspond to nodes that are
endpoints of e. Thus C must contain at least one
endpoint of e.

Solutions
Algorithm 1: (in the case of uniform cost)
1: C = empty
2: while U is not empty
3: pick a set Si such that Si covers the most
elements in U
4: remove the new covered elements from U
5: C = C union Si
6: return C

Solutions (cont)
 In the case of non-uniform cost
 Similar method. At each iteration, instead of picking a
set Si such that Si covers the most uncovered elements,
we will pick a set Si whose cost-effectiveness α is
smallest, where α is defined as:
 Questions: Why choose smallest α? Why define α as
above
||/)( USSc ii

Solutions (cont)
Algorithm 2: (in the case of non-uniform cost)
1: C = empty
2: while U is not empty
3: pick a set Si such that Si has the smallest α
4: for each new covered elements e in U
5: set price(e) = α
6: remove the new covered elements from U
7: C = C union Si
8: return C

Approximation Ratio Analysis
Let ek, k = 1…n, be the elements of U in the order in which
they were covered by Alg 2. We have:
 Lemma 1:
 Proof: Let Uj denote remaining set U at iteration j. That
is, Uj contains all the uncovered elements at iteration j.
At any iteration, the leftover sets of the optimal solution
can cover the remaining elements at a cost of at most
opt. (Why?)
solutionoptimaltheofcosttheiswhere
)1/()(price},,...,1{eachFor
opt
knoptenk k

Proof of Lemma 1 (cont)
Thus, among these leftover sets, there must exist
one set where its α value is at most opt/|Uj|. Let
ek be the element covered at this iteration, |Uj| ≥
n – k + 1. Since ek was covered by the most
cost-effective set in this iteration, we have:
price(ek) ≤ opt/|Uj| ≤ opt/(n-k+1)

Approximation Ratio
 Theorem 1: The set cover obtained from Alg 2
(also Alg 1) has a factor of Hn where Hn is a
harmonic series Hn = 1 + 1/2 + … + 1/n
 Proof: It follows directly from Lemma 1

Examples of NP-complete problems
 Independent set
- independent set: a set of vertices in the graph with no
edges between each pair of nodes.
- given a graph G=(V,E), find the largest independent set
- reduction from vertex cover:
smallest vertex cover S
largest independent set
V/S

Independent Set
 Independent set
- if G has a vertex cover S, then V S is an independentset
proof: considertwo nodes in V S:
if there is an edge connectingthem, then one of them must be in
S, which means one of themis not in V S
- if G has an independentset I, then V I is a vertex cover
proof: consideroneedge in G:
if it is not covered by any nodein V I, then its two end vertices
must be both in I, which means I is not an independent set

Summary of some NPc problems
SAT
3SAT
Vertex cover
Independent
set
Set
cover
Graph coloring
Maximum
clique size
Minimum
Steiner
tree
Hamiltonian
cycle
Maximum cut
find more NP-complete problems in
http://en.wikipedia.org/wiki/List_of_NP-complete_problems

Lecture26

More Related Content

What's hot (20)

Viewers also liked (20)

Similar to Lecture26 (20)

More from Dr Sandeep Kumar Poonia (17)

Recently uploaded (20)

Lecture26