1. Algorithms & LPs for k-Edge Connected Spanning Subgraphs
David Pritchard
Zinal Workshop, Jan 20 ‘09

2. k-Edge Connected Graph
k edge-disjoint paths between every u, v
at least k edges leave S, for all ∅ ≠ S ⊊ V
even if (k-1) edges fail, G is still connected
|δ(S)| ≥ k S

3. k-ECSS & k-ECSM Optimization Problems
k-edge connected spanning subgraph problem (k-ECSS): given an initial graph (maybe with edge costs), find k-edge connected subgraph including all vertices, w/ |E| (or cost) minimal
k-ecs multisubgraph problem (k-ECSM): can buy as many copies as you like of any edge G
3-edge-connected multisubgraph of G, |E|=9

4. Approximation Algorithms
k-ECSS and k-ECSM are NP-hard for k ≥ 2 For k-ECS/M or any NP-hard minimization problem, one notion of a good heuristic is a polytime “α-approximation algorithm”: always produces a feasible solution (e.g., k-edge connected subgraph) whose output cost is at most α times optimal Also call α the approximation ratio/factor

5. Approximability Any problem has either
no constant-factor approximation algorithm
Ǝ constant α ≥ 1: α-approx alg exists, but no (α-ε)-approx exists for any ε>0 [α=1: “in P”]
Ǝ constant α ≥ 1: (α+ε)-approx alg exists for any ε>0, but no α-approx [α=1: “apx scheme”]

6. Question
What is the approximability of the k-edge-connected spanning subgraph and k-edge-connected spanning multisubgraph problems?
It will depend on k
Exact approximability not known but we’ll determine rough dependence on k
Also look at important unit-cost special case

7. History of k-ECSS
2-ECSS is NP-hard, has a 2-apx alg [FJ 80s] & is APX-hard (has no approximation scheme if P ≠ NP) even for unit costs [F 97]
k-ECSS has a 2-apx alg [KV 92]
unit cost k-ECSS: has a (1+2/k)-apx [CT96] & Ǝ tiny ε>0, for all k>2, no (1+ε/k)-apx [G+05]
Gets easier to approximate as k increases!
How do k-ECSS, k-ECSM depend on k?

8. Main Course
We prove k-ECSS with general costs does not have approximation ratio tending to 1 as k tends to infinity
We conjecture that k-ECSM with general costs does have approximation ratio tending to 1 as k tends to infinity

9. Part 1: Approximation Hardness

10. An Initial Observation
For the k-ESCM (multisubgraph) problem, we may assume edge costs are metric, i.e.
cost(uv) ≤ cost(uw) + cost(wv)
since replacing uv with uw, wv maintains k-EC u v S w

11. What’s Hard About Hardness?
A 2-VCSS is a 2-ECSS is a 2-ECSM.
For metric costs, can split-off conversely, e.g.
All of these are APX-hard [via {1,2}-TSP]
vertex-connected
2-ECSM
2-ECSS
2-VCSS

12. What’s Hard About Hardness?
1+ε hardness for 2-VCSS implies 1+ε hardness for k-VCSS, for all k ≥ 2
But this approach fails for k-ECSS, k-ECSM G
G, a hard instance for 2-VCSS
zero-cost edges to V(G)
Instance for 3-VCSS with same hardness

13. Hardness of k-ECSS (slide 1/2) Ǝ ε>0, ∀ k≥2, no 1+ε-apx if P ≠ NP
Reduce APX-hard TreeCoverByPaths to k-ECSS
Input: a tree T, collection X of paths in T
A subcollection Y of X is a cover if the union of {E(p) | p in Y} equals E(T)
Goal: min-size subcollection of X that is a cover
size-2 cover

14. Hardness of k-ECSS (slide 2/2) Ǝ ε>0, ∀ k≥2, no 1+ε-apx if P ≠ NP
Replace each edge e of T by k-1 zero-cost parallel edges; replace each path p in X by a unit-cost edge connecting endpoints of p
… min |X| to cover T = k-ECSS optimum.
0 × (k-1)
0 × (k-1)
0 × (k-1)
0 × (k-1)
0 × (k-1)
0 × (k-1) 1 1 1 1

15. Part 2: Linear Programs

16. From Hardness to Approximability
Conjecture [P.] For some constant C, there is a (1+C/k)-approximation algorithm for k-ECSM.
For k ≤ 2 it is known to hold with C=1.
Next: definition of LP-relative, motivation
an LP-relative

17. LP Relaxation Introduce variables xe ≥ 0 for all edges e of G.
LP relaxation of k-ECSM problem:
min ∑ xecost(e) s.t. x(δ(S)) ≥ k for all ∅ ≠ S ⊊ V
0.4
∑e in δ(S) xe ≥ k S
1.2
1.4

18. LP-Relative a new name for an old concept
The LP is a linear relaxation of k-ECSM so LP-OPT ≤ OPT (optimal k-ECSM cost)
Definition: an LP-relative α-approximation algorithm has output cost ALG ≤ α⋅LP-OPT
(Stronger than “α-approx” — ALG ≤ α·OPT) +
LP-OPT
OPT
ALG
(integrality gap)

19. Motivation Similar results for related problems are true
LP-relative (1+C/k)-approx for k-ECSM would imply a (1+C/k)-approx for subset k-ECSM
Subset k-ECSM: given a graph with some vertices required, find a graph with edge-connectivity ≥ k between all required vertices (k=1: Steiner tree)
To show this, use parsimonious property [GB93] of the LP: there is an optimal fractional solution that doesn’t use non-required vertices

20. A Combinatorial Approach?
“Ǝ constant C so that for every A, B > 0, every (A + B + C)-edge-connected graph contains a disjoint A-ECSM and B-ECSM” would imply the conjecture. Cousins: Ǝk, each k-strongly connected digraph has 2 disjoint strongly connected subdigraphs Ǝk, every k-edge connected hypergraph contains 2 disjoint connected hypergraphs
OPEN [B-J Y 01]
FALSE [B-J T 03]

21. More About the LP
LP relaxation of k-ECSM equals (up to scaling) the Held-Karp TSP relaxation, and the undirected Steiner tree LP relaxation.
LP and generalizations are well-studied
Basic (“extreme”) solutions have few nonzeroes
Extreme solution properties are often key to algorithmic results (e.g., [GGTW 05])
How “unstructured” can extreme points get?

22. Extremely Extreme Extreme Point
Edge values of the form Fibi/Fib|V|/2 and 1 - Fibi/Fib|V|/2 (exponentially small in |V|)
Maximum degree |V|/2

23. Structural LP Property [CFN]
The LP has 2|V|-1 constraints, |V|2 vars, but
Ǝ a basic solution which is optimal (all LPs);
we can find it in polytime (reformulation);
every basic solution has ≤ 2|V|-3 nonzero coordinates (laminar family of constraints).

24. 1+O(1/k) Algorithm for Unit-Cost k-ECSM [GGTW]
Solve LP to get a basic optimal solution x*
Round every value in x* up to the next highest integer and return the corresponding multigraph
(k=4)

25. Analysis
Optimal k-ECSM has degree k or more at each vertex, hence at least k|V|/2 edges
The fractional LP solution x* has value (fractional edge count) k|V|/2
There are at most 2|V|-3 nonzero coordinates
Rounding up increases cost by at most 2|V|-3
ALG/OPT ≤ (k|V|/2 + 2|V|-3)/(k|V|/2) < 1 + 4/k

26. Open Questions
Resolve the conjecture! What’s the minimum f(A, B) so that any f(A, B)-edge-connected (multi)graph contains disjoint A- and B-edge-connected subgraphs? Is there a constant k such that every k-strongly connected digraph has 2 disjoint strongly connected subdigraphs?

27. Thanks for Attending!

28. Small Extreme Examples
n=7, Δ=4
n=6, denom=2
n=8, denom=3
n=9, Δ=5
n=9, denom=4
n=10, denom=Δ=5

29. Previously Known Constructions
[BP]: minimum nonzero value of x* can be ~1/|V|
[C]: max degree can be ~|V|1/2