Approximation Algorithms

From MyInfoRepo
Jump to navigation Jump to search

Approximation algorithms provide a scheme.. c-gap... Examples..


Overview

inapprox overview For optimization problems, we want a solution that has an associated (min/max) cost. For a set of given instances (or bounded instances), we approximate the best non negative solution.

OPT(x) = optimal cost of, or solution itself for an instance. All solutions have polynomial length; NPO = {optimization problems}

Converting an instance of NPO:

min: is OPT(x) <= val? max: is OPT(x) >= val?

These problems are in NP, since solutions are polynomially solvable.

approximation: ALG is a c-approx. if min: cost(ALG(x))/OPT <= c (c >= 1) max: OPT/cost(ALG(x)) <= c (c >= 1)

PTAS poly-time approx scheme

The alg takes an additional input eps > 0, producing a 1+eps solution. Poly-time for fixed eps. PTAS = {NPO problems with alg} F-APX = {NPO with poly f(n)-approx alg for some f in F} APX = O(1)-APX Log-APX = Ologn-APX PTAS is a strict subset of APX

approx reductions

approximation preserving reductions

go from instance A : x -> B : x' with function f x' = f(x) If we find a solution for B y' to x', convert it back to A, with function g. y=g(x,y')

PTAS reduction

Finding a 1+eps approximation for A, by finding a 1+d(eps) for B. f and g depend on eps. if B in PTAS* then A in PTAS*, if A not in PTAS* then B not in PTAS*.

  • also holds for APX.

AP-reduction: d(eps) = O(e). if B has an O(f)-apx then A O(f)-apx. strict reduction: d(e) = e APX-hard = PTAS-reduction from any problem in APX =/= PTAS. Otherwise all problems would have a PTAS.

L-reduction

- OPTb(x') = O(OPTa(x))
- |costa(y) - costa(x)|, absolute error is O(|costb(y') - costb(x')|)

d(eps) = A*B*eps

min: costa(y)/OPTa(x) < OPTa(x)+B(costb(y)-OPTb(x) so costa > OPT


Approximation classes

PO < FPTAS < PTAS < APX < NPO This holds iff P = NP. Relative inapproximability, problems in APX where the apx-quality cannot be improved beyond a constant. Absolute inapproximability, problems that have no c-apx alg for c < inf. (problems in NPO - APX)

gap introducing and gap amplification can be used Gap introducing can be used to establish both rel and abs inapprox. Amplification can show absolute inapprox.

Gap introducing reductions

Suppose we have an optimization problem B. The objective is to minimize the cost for B. Here we focus on the possibility of finding a approx alg, with a certain quality. Using an NP-hard decision problem A, which has yes/no instances. Reducing no instances to problem B, where OPT > b, and yes instances to OPT <= a, using poly function f. There is an (a,b] gap between the two instances. We prove that such a reduction exists if these properties are satisfied. The reduction leads to an approximation algorithm of at best b/a-apx quality. There does not exist an apx with a better approximation than b/a.

proof: If such an algorithm, lets call it ALGb/a would exist, for any input x on problem a, run f(x). This creates an instance of B. Now, running ALGb/a will produce a result x', which is either smaller than b, resulting in a yes, otherwise it will result in a no. To show that the ALG is valid we show: yes from problem A -> f(x), the ALGb/a(f(x)) produces < b/a*OPT < b, so the output produces yes. no from problem A -> f(x), the ALGb/a(f(x)) produces > b/a*OPT > b -> no output

gap introducing proof for minbinpacking

NP-complete problem: partitioning, partition into two sets of equal weight. Let f be the reduction from partitioning to binpacking: on input x = (A,g) in Partitioning, let f(x) = (S,w) in MinBinPacking where - S = A - weight function: w = g(a) / C, where C is half of the weight C = 0.5*Totalweight. Separate outcomes for approximation non-existance: It is clear that if x = (A,g) = yes-instance, 2 bins suffice, therefore OPT(f(x)) < 2. x = no-instance, when 3 bins are necessary. Therefore, OPT(f(x)) > (3-eps) for eps in (0,1]