# Short Cycles make W-hard problems hard: FPT algorithms for W-hard Problems in Graphs with no short Cycles

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1 Short Cycles make W-hard problems hard: FPT algorithms for W-hard Problems in Graphs with no short Cycles Venkatesh Raman and Saket Saurabh The Institute of Mathematical Sciences, Chennai {vraman Abstract. We show that several problems that are hard for various parameterized complexity classes on general graphs, become fixed parameter tractable on graphs with no small cycles. More specifically, we give fixed parameter tractable algorithms for Dominating Set, t-vertex Cover (where we need to cover at least t edges) and several of their variants on graphs with girth at least five. These problems are known to be W [i]-hard for some i 1 in general graphs. We also show that the Dominating Set problem is W [2]-hard for bipartite graphs and hence for triangle free graphs. In the case of Independent Set and several of its variants, we show these problems to be fixed parameter tractable even in triangle free graphs. In contrast, we show that the Dense Subgraph problem where one is interested in finding an induced subgraph on k vertices having at least l edges, paramaterized by k, is W [1]-hard even on graphs with girth at least six. Finally, we give an O(log p) ratio approximation algorithm for the Dominating Set problem for graphs with girth at least 5, where p is the size of an optimum dominating set of the graph. This improves the previous O(log n) factor approximation algorithm for the problem, where n is the number of vertices of the input graph. 1 Introduction Parameterized complexity is a practical approach to deal with intractable computational problems having some small parameters. For decision problems with input size n, and a parameter k (which typically, and in all the A preliminary version of this paper appeared as Triangles, 4-cycles and Parameterized (In-)Tractability in the proceedings of 10th Scandinavian Workshop on Algorithm Theory (SWAT 2006) [24].

2 2 problems we consider in this paper, is the solution size), the goal here is to design an algorithm with runtime f(k)n O(1) where f is a function of k alone, as contrasted with a trivial n k+o(1) algorithm. Problems having such an algorithm is said to be fixed parameter tractable (FPT), and such algorithms are practical when small parameters cover practical ranges. The book by Downey and Fellows [8] provides a good introduction to the topic of parameterized complexity. For recent developments see the books by Flum and Grohe [14] and Niedermeier [17]. There is a hierarchy of intractable parameterized problem classes above FPT, the main ones are: F P T M[1] W [1] M[2] W [2] W [P ] XP. The principal analogue of the classical intractability class NP is W [1]. A convenient source of W [1]-hardness reductions is provided by the result that Independent Set is complete for W [1] [8]. Other highlights of the theory include that Dominating Set, by contrast, is complete for W [2] [8]. Surprisingly we show that these problems and several of their variants that are known to be hard in the W-hierarchy, are fixed parameter tractable on graphs that have no short cycles more specifically on graphs with girth at least five. These problems are known to be NP-complete on such graphs as well [4, 5]. We also look at the Set Cover problem where the size of the intersection of any pair of sets is bounded by a fixed constant. While the general version of Set Cover is known to be W [2]-complete, we prove this special version fixed parameter tractable. Most of our algorithms are based on the method of kernelization. The main idea of kernelization is to replace a given instance (I, k) by a simpler instance (I, k ) using some data reduction rules in polynomial time such that (I, k) is a yes instance if and only if (I, k ) is a yes instance and I is bounded by a function of k alone. The reduced instance is called kernel for the problem. For most of our problems we give polynomial sized kernel in polynomial time. 1.1 Organization of the Rest of the Paper In Section 2, we look at the Dominating Set problem and show that the problem is W [2]-complete even in bipartite graphs and split graphs (a graph in which the vertices can be partitioned into a clique and an independent set). Though variations of Dominating Set like Red-Blue Dominating Set [10] and Constrained Dominating Set [16] have

3 been studied before and shown to be W[2]-complete, to the best of our knowledge the standard Dominating Set problem (which we consider here) in bipartite graphs has not been studied before. Our observation means that the dominating set problem is W [2]-complete in triangle free graphs. Then we show that the problem is FPT if the input graph has girth at least 5. It turns out that this result can be generalized to several variants of the Dominating Set problem on graphs with girth at least five. In Section 3, we look at the Set Cover problem for which Dominating Set is a special instance. Set Cover problem is known to be W [2]-complete [8]. Here we show that if the set cover instance satisfies the property that the intersection of any pair of its sets is bounded by a fixed constant then the problem is fixed parameter tractable. In Section 4, we look at t-vertex Cover and t-dominating Set problems. These are generalizations of Vertex Cover and Dominating Set problems. In the t-vertex Cover problem, we are interested in finding a set of at most k vertices covering at least t edges and in the t-dominating Set problem the objective is to find a set of at most k vertices that dominates at least t vertices. Both these problems have been parameterized in two different ways: by k alone and by both k and t. Both these problems are fixed parameter tractable when parameterized by both k and t. Bläser [6] gave O(2 O(t) n O(1) ) algorithm for both the problems using color coding technique. Guo et. al. [18] have shown that t-vertex Cover is W [1]-complete when parameterized by k alone. It is easy to see that the t-dominating Set is W [2]-complete by a reduction from Dominating set when parameterized by k alone. We show that both these problems are fixed parameter tractable in graphs with girth at least five, when parameterized by k alone. In Section 5, we look at the Independent Set problem and several of its variants. We show that these problems are fixed parameter tractable in triangle free graphs while they are W[1]-complete in general graphs. In contrast to our results in earlier sections, in Section 6, we exhibit a problem that is W [1]-hard in graphs with no small cycles. This is the Dense Subgraph problem [22]. Here, given a graph G = (V, E) and positive integers k and l, the problem is to determine whether there exists a set of at most k vertices C V such that the induced subgraph on C has at least l edges; here k is the parameter. In Section 7, we deviate and look at the approximability result of the Dominating Set problem. We conclude that the Dominating Set problem is as hard to approximate in bipartite graphs as in general undi- 3

4 4 rected graphs. We also give an approximation algorithm of factor O(log p) for the Dominating Set problem if the input graph has girth at least 5, where p is the size of an optimum dominating set of the input graph. This improves the previously known approximation algorithm of factor O(log n), where n is the number of vertices in the input graph. Section 8 gives some concluding remarks and open problems. We assume that all our graphs are simple and undirected. Given a graph G = (V, E), n represents number of vertices, and m represents the number of edges. For a subset V V, by G[V ] we mean the subgraph of G induced on V. By N(u) we represent all vertices (excluding u) that are adjacent to u, and by N[u], we refer to N(u) {u}. Similarly, for a subset D V, we define N[D] = v D N[v]. By the girth of a graph, we mean the length of the shortest cycle in the graph. We say that a graph is a G i graph if the girth of the graph is at least i. A vertex is said to dominate all its neighbors. 2 Dominating Set and its Variants In this section we look at the Dominating Set problem and its variants. Dominating Set: Given a graph G = (V, E) and an integer k 0, determine whether there exists a set D V, of size at most k, such that for every vertex u V, N[u] D. We say that the set D dominates the vertices of G. We first show that Dominating Set problem is W[2]-complete in bipartite graphs and split graphs by a reduction from the same problem in general undirected graphs. Then we give a fixed parameter tractable algorithm for the problem in graphs with girth at least Dominating Set in Bipartite and Split Graphs Theorem 1. Dominating Set problem is W [2]-complete in bipartite graphs. Proof. We prove this by giving a reduction from the Dominating Set problem in general undirected graphs. Given an instance (G = (V, E), k) of Dominating Set, we construct a bipartite graph H = (V, E ). Let z 1 and z 2 be two new vertices (not in V ). Now V = V 1 V 2 where V 1 = {u 1 u V } {z 1 } and V 2 = {u 2 u V } {z 2 }. If there is an edge (u, v) in E then we draw the edges (u 1, v 2 ) and (v 1, u 2 ). We also draw

5 edges of the form (u 1, u 2 ) for every u V. Finally, we add an edge from every vertex in V 1 to z 2. This completes the construction of H. We show that G has a dominating set of size k if and only if H has a dominating set of size k + 1. Let D be a dominating set of size k in G. Then clearly D = {u 1 u D} {z 2 } is a dominating set of size k + 1 in H. Conversely, let K be a dominating set in H of size k + 1. Observe that either z 1 or z 2 must be part of K as z 2 is the unique neighbor of z 1. Without loss of generality, we can assume that z 2 K, as otherwise we could delete z 1 and include z 2 in K and still have a dominating set of size at most k + 1 in H. Now take D = {u u V, u 1 or u 2 K}. Clearly D is of size k. We show that D is a dominating set in G. For any u / D, u 2 / K and hence there exists some v 1 K such that v 1 dominates u 2 in H. But this implies v D and (v, u) E, which shows that v dominates u. This proves that D is a dominating set of size k for G and establishes the theorem. Since every bipartite graph is also triangle free, we have the following corollary. Corollary 1. Dominating Set problem is W [2]-complete in triangle free graphs. Theorem 2. Dominating Set problem is W [2]-complete in split graphs. Proof. We again prove this by giving a reduction from the Dominating Set problem in general undirected graphs. Given an instance (G = (V, E), k) of Dominating Set, we construct a split graph H = (V, E ). We create two copies of V namely V 1 = {u 1 u V } and V 2 = {u 2 u V }. If there is an edge (u, v) in E then we draw the edges (u 1, v 2 ) and (v 1, u 2 ). We also draw edges of the form (u 1, u 2 ) for every u V. Now we make H[V 1 ] a complete graph by adding all arcs of the form (u 1, v 1 ) for every pair of vertex u 1, v 1 V 1. This completes the construction of H. It is easy to see that H is a split graph with H[V 1 ] as a clique and H[V 2 ] as an independent set. As in the proof of Theorem 1, it is easy to see that G has a dominating set of size k if and only if H has a dominating set of size k. An undirected graph is chordal if every cycle of length greater than three possesses a chord, that is, an edge joining two nonconsecutive vertices of the cycle. It is well known that every split graph is also a chordal graph and hence Theorem 2 implies that the Dominating Set is W [2]-complete in chordal graphs. As a corollary of Theorem 2 we get the following. 5

6 6 Corollary 2. Dominating Set problem is W [2]-complete in chordal graphs. 2.2 FPT Algorithm for Dominating Set in G 5 Graphs We give a fixed parameter tractable algorithm for the Dominating Set problem in graphs with girth at least 5 (G 5 graphs) and also observe that various other W -hard problems become tractable for G 5 graphs. Our algorithm follows a branching strategy where at every iteration we find a vertex that needs to be included in the Dominating Set which we are trying to construct. Once a vertex is included, we can at best delete that vertex. Though the neighbors of the vertex are dominated, we can not remove these vertices from further consideration as they can be useful to dominate other vertices. Hence we resort to a coloring scheme for the vertices, similar to the one suggested by Alber et al. in [2, 3]. At any point of time of the algorithm, the vertices are colored as below: 1. Red - The vertex is included in the dominating set D which we are trying to construct. 2. White - The vertex is not in the set D, but it is dominated by some vertex in D. 3. Black - The vertex is not dominated by any vertex of D. Now we define the dominating set problem on the graph with vertices colored with White, Black or Red as above. We call a graph colored red, white and black as above, as a rwb-graph. RWB-Dominating Set: Let G be a G 5 graph (graph with girth at least 5) with vertices colored with Red, White or Black satisfying the following conditions, and let k be a positive integer parameter. Let R, W and B be the set of vertices colored red, white and black respectively. 1. Every white vertex is a neighbor of a red vertex. 2. Black vertices have no red neighbors. 3. R k Does G have at most k R vertices that dominate all the black vertices? It is easy to verify that if we start with a general G 5 graph with all vertices colored black, and color all vertices we want to include in the dominating set as red, and their neighbors as white, the graph we obtain

7 at every intermediate step is a rwb-graph, and the problem we will have at the intermediate steps is the RWB-Dominating Set problem. The following lemma essentially shows that if the rwb-graph has a black or white vertex dominating more than k black vertices, then such a vertex must be part of every solution of size at most k, if one exists. Lemma 1. Let (G = (R W B, E), k) be an instance of the RWB- Dominating Set problem where G is a G 5 graph and k a positive integer. Let v be a black or white vertex with more than k R black neighbors. Then if G has a set of size at most k R that dominates all black vertices, then v must be part of every such set. Proof. Let D be a set of size k R that dominates all black vertices in G, and suppose v / D. Let X be the set of black neighbors of v which are not in D and Y be the set of black neighbors of v in D. So X + Y > k R. Observe that for every v x X we have a neighbor u x D which is not in Y (otherwise v, v x, u x is a 3 length cycle). Similarly, for x, y X, x y u x u y. Otherwise v, x, u x, y will form a cycle of length 4. This means that D X + Y > k R which is a contradiction. Given a rwb-graph, Lemma 1 suggests the following simple reduction rule. (R1) If there is a white or a black vertex v having more than k R black neighbors, then color v red and color its neighbors white. Our goal now is to pick enough white or black vertices to dominate the black vertices. So the following reduction rules are obviously justified. (R2) If a white vertex is not adjacent to a black vertex, delete the white vertex. (R3) If there is an edge between two white vertices, delete the edge. (R4) If R > k, then stop and return NO. The following lemma follows from Lemma 1. Lemma 2. Let G = (R W B, E) be an instance of RWB-Dominating Set and let G = (R W B, E ) be the reduced instance after applying rules (R1) to (R4) once. Let k be an integer parameter. Then G is a yes instance if and only if G is a yes instance. That is G has a set of size at most k R dominating all vertices in B if and only G has a set of size at most k R dominating all vertices in B. 7

8 8 Let G be an instance of RWB-Dominating Set and let G be the reduced instance after applying the reduction rules (R1) (R4) until no longer possible. Then we show that if G is a yes instance (and hence G is a yes instance), the number of vertices in G is bounded by polynomial in k. More precisely we show the following lemma. Lemma 3. Let (G, k) be a yes instance of RWB-Dominating Set and (G, k ) be the reduced instance of (G, k) after applying the rules (R1) R(4) until no longer possible. Then, the number of vertices in G is O(k 3 ), that is, a kernel of size at most O(k 3 ) can be obtained for RWB- Dominating Set. Proof. Let R, B and W be the set of vertices colored red, black and white respectively in G. We argue that each of R, B and W is bounded by a function of k. Because of (R4) (and the fact that G is a yes instance), R k. Because of (R1), every vertex colored white or black has at most k R black neighbors. Also we know that no red vertex has a black neighbor. Since G is a yes instance, there are at most k (k R to be more precise) black or white vertices dominating all black vertices. Since each of them can dominate at most k black vertices, we conclude that B can be at most k 2. We argue that W k 3. Towards this end, we just show that every black vertex has at most k white neighbors. Since B k 2, and every white vertex is adjacent to some black neighbor (because of (R2) and (R3)), the conclusion will follow. Note that every white vertex has a red neighbor. Observe that the white neighbors of any black vertex (any vertex for that matter) will have all distinct red neighbors. I.e. if w 1 and w 2 are white neighbors of a black vertex b, then there is no overlap between the red neighbors of w 1 and the red neighbors of w 2. This is because if w 1 and w 2 have a common red neighbor r, then we will have a 4-cycle b, w 1, r, w 2, b. Since R k, it follows that a black vertex can have at most k white neighbors. This proves the required claim. Thus we have the following theorem. Theorem 3. The RWB-Dominating Set problem can be solved in O(k k+o(1) + n O(1) ) time for G 5 graphs. Proof. It is easy to see that the reduction rules (R1) to (R4) take polynomial time to execute. When none of these rules can be executed, by

9 Lemma 3, we have that the number of vertices in the resulting graph is O(k 3 ), and each vertex has at most k black neighbors. We can just try all possible subsets of size at most k of the vertex set of the reduced graph, to see whether that subset dominates all the black vertices. If any of them does, then we say YES and NO otherwise. This will take O(k 3k+O(1) ) time. Alternatively, we can apply a branching technique on the black vertices, by selecting a black vertex or any of its neighbors in the dominating set. More precisely, let v be a black vertex. Then we branch on N[v] by including w N[v] in the possible dominating set D we are constructing and look for a solution of size k 1 in G {w} where w is colored red and all its neighbors are colored white for every w N[v]. This results in an O((k + 1) k+o(1) ) time algorithm. Now to solve the general Dominating Set problem in G 5 graphs, simply color all vertices black and solve the resulting RWB-Dominating Set problem using Theorem 3. Thus we have Theorem 4. Parameterized Dominating Set problem can be solved in O(k k+o(1) + n O(1) ) time for G 5 graphs. Parameterized version of Connected Dominating Set (where one is interested in dominating set which is connected) or Independent Dominating Set (where one is interested in dominating set which is independent) are also known to be W[2]-complete [8]. Since the reduction rules (R1)-(R4) apply for any dominating set, using Lemma 3 we can obtain a kernel of size at most O(k 3 ) for both these problems. For the Independent Dominating Set problem we also check that R remains an independent set when we add a vertex to it while applying reduction rule (R1), else we return NO. Furthermore in the proof of the Theorem 3, we try all possible subsets of size at most k and look for a connected or independent dominating set, as required. This results in the following corollary. Corollary 3. Parameterized Connected Dominating Set and Independent Dominating Set problems can be solved in O(k 3k+O(1) +n O(1) ) time for G 5 graphs. A number of other variants of dominating set problem which are W[2]- hard can be shown to be fixed parameter tractable in a similar way for G 5 graphs though not using kernelization. We give necessary details for a few of them in the next subsections. 9

10 Red-Blue Dominating Set and Constraint Bipartite Dominating Set In this section we give FPT algorithms for Red-Blue Dominating Set and Constrained Bipartite Dominating Set problems for G 5 graphs. We first give an algorithm for Red-Blue Dominating Set problem which is defined as follows. Red-Blue Dominating Set [10]: Given a bipartite graph G = (V, E) with V bipartitioned as V red V blue and a positive integer k. Does there exist a subset D V red with D k and V blue N(D). Theorem 5. Red-Blue Dominating Set is FPT for G 5 graphs. Proof. Any two vertices in V red have at most one common neighbor in V blue as otherwise there will be a four cycle in G. Hence, the following reduction rule is justified. (R1 ) if x V red has degree more than k then include x D. The correctness of (R1 ) follows from the fact that if we do not select x in D then we need more than k vertices from V red to dominate N(x) as any vertex y V red, y x, can dominate at most one vertex of N(x). Hence after exhaustively applying reduction rule (R1 ) if the size of D is more than k we answer NO. Remove N[D] from G, i. e., set V red = V red \D and V blue = V blue \N(D). Now the degree of every vertex in V red is at most k and we are looking for a set of size at most k D in V red such that it dominates all the vertices of V blue. Since every vertex in V red has degree at most k, the size of the set V blue is bounded above by k 2 ((k S )k to be precise) else the answer is NO. We can not bound the size of the set V red anymore, as we do not have any bound on the degree of the vertices in V blue. So to find the desired dominating set in V red (dominating all the vertices in V blue ) we do as follows: For all partitions P of V blue into at most k D parts, say P = {P 1, P 2,, P j }, 1 j k D, for each P i, 1 i j check whether there exists a vertex u i V red such that P i N(u i ). Call the partition P valid if for all 1 i j, there exists u i V blue such that P i N(u i ) and the set {u i 1 i j} is called the witness set. If any partition P is valid then return YES with the corresponding witness set else return NO.

11 It is easy to see that there exits a subset of V red of size at most k D dominating all vertices of V blue if and only if there exists a valid partition. Number of ways in which n indistinguishable objects can be partitioned into r ways is ( ) n+r 1 r 1 [20]. Hence the total number of partitions P considered for our case is upper bounded by k D i=1 ( k 2 + i 1 i 1 Since the total number of partitions is upper bounded by O(k 2k+O(1) ), the result that Red-Blue Dominating Set is FPT for G 5 graphs follows. Next we study Constraint Bipartite Dominating Set problem which is defined as follows. Constraint Bipartite Dominating Set (CBDS) [16]: Given a bipartite graph G = (V, E) with V partitioned as V 1 V 2 and positive integers k 1 and k 2. Does there exist subsets D 1 V 1 and D 2 V 2 with D 1 k 1 and D 2 k 2 such that V 2 N(D 1 ) and V 1 N(D 2 ). Theorem 6. Parameterized Constraint Bipartite Dominating Set is FPT for G 5 graphs. Proof. To solve this problem we just need to solve two instances of Red- Blue Dominating Set problem. The instances of Red-Blue Dominating Set problem we solve are: ). 1. V red = V 1, V blue = V 2 and parameter is k 1 ; and 2. V red = V 2, V blue = V 1 and parameter is k 2. We return YES for CBDS problem if both the instances return YES and as D 1 the red-blue dominating set returned by instance 1 and as D 2 the red-blue dominating set returned by instance 2. If either of the instances of Red-Blue Dominating Set problem returns NO, then we return NO for the CBDS problem Threshold Dominating Set This problem generalizes Dominating Set and is formally defined as follows.

12 12 Threshold Dominating Set (TDS) [7]: Given a graph G = (V, E) and positive integers k and r. Is there a set of at most k vertices V V such that for every vertex u V, N[u] contains at least r elements of V? Theorem 7. Threshold Dominating Set parameterized by k is FPT for G 5 graphs. Proof. First we observe that if k < r, then the answer is NO. We assume that r log n, as otherwise k log n and we have a kernel of size at most 2 k. Now we can solve the problem by checking all subsets of size at most k for the desired threshold dominating set. Our algorithm is again based on the following simple reduction rule whose correctness follows from Lemma 1. (R1 ) if x V has degree more than k then include x V. So basically we select all the vertices of degree more than k of V in V and hence if the size of V is more than k then we answer NO. Next we assign a color to all the vertices. We assign white color to all the vertices (including vertices in V ) which have enough (at least r) neighbors in V and black to the rest. Let B and W, as usual, represent the set of black and white vertices respectively and set B = B \ V and W = W \ V. Apply reduction rules (R2) and (R3) of Lemma 2 exhaustively. The rule (R3) makes G[W ] an independent set. Now the problem reduces to finding a set S of size at most k V in V \ V such that V S is a desired threshold dominating set for G, in particular for the vertices of B. Since every vertex in V \ V has degree at most k and we are looking for S of size at most k in V \V, the size of B is bounded above by k 2, as otherwise we answer NO. Now what we have is a generalized version of Threshold Dominating Set problem where we have a set of j k 2 black vertices B = {u 1,, u j }, each with a positive integer r i (r i = r N[v i ] V ), 1 i j. We are looking for a set S (W B ) of size at most k V such that for every u i B, N(u i ) S r i in G where the vertex set of G is V (G ) = B W and the edge set of G is E(G ) = {(u, v) E u W, v B or u B, v B}. To solve this generalized version of Threshold Dominating Set problem, we need to generalize our partition arguments used in the Theorem 5 suitably. The major differences are that G is no more bipartite and that there are vertices which need more than 1 (possibly r) vertices in the desired threshold dominating set. To overcome this difficulty, we

13 make a multiset M from B by having r i copies for each vertex u i B. Clearly the size of M is bounded above by rk 2. Now if we apply the partition idea of Theorem 5 it is possible that the same vertex may dominate multiple copies of the same vertex. To deal with this call a partition P = {P 1, P 2,, P α } valid if (a) there exists a subset S B W forming a system of distinct representatives; that is for all 1 i α, there exists a distinct u i S such that P i N(u i ) and (b) each P i contains at most one copy of any vertex of B. The set S is a witness set. So to find the desired threshold dominating set in B W we proceed as follows. For all partitions P of M in at most k V parts, say P = {P 1, P 2,, P α }, 1 α k V, we check whether P is a valid partition. If any partition P is valid then return YES with the corresponding witness set else return NO. For a fixed partition P = {P 1, P 2,, P α }, we can do the validity testing and find a corresponding witness set in polynomial time as follows. Testing for duplicate copies in P i s are easy. For the other part we first define the set I i = {u (B W ) P i N G (u)}, where N G (u) denotes the neighbors of u in G. Now we make the bipartite incidence graph for the sets {I 1,, I α }, that is a bipartite graph G = (X Y, E ), where X has a vertex x i for every set I i and Y = α l=1 I l and there is an edge between (x i, u) if u I i. Now finding a valid system of distinct representatives reduces to finding a maximum bipartite matching in G saturating X, for which there is a classical polynomial time algorithm of Edmonds [12]. The total number of partitions P considered for our case is upper bounded by k V i=1 ( rk 2 + i 1 i 1 which is at most O((rk 2 ) k+o(1) ). Since r log n and (log n) k n + (2k log k) k for all n and k n, we have the desired result that Threshold Dominating Set problem is FPT for G 5 graphs. 3 Set Cover with Bounded Intersection among Sets Dominating Set problem is well known to be a special instance of the Set Cover problem defined below. ), 13

14 14 Set Cover: Given a base set (or universe) U = {s 1, s 2,, s n }, a collection S = {S 1, S 2,, S m } of subsets of U (S i U, 1 i m) such that m i=1 S i = U and a positive integers k, does there exist a sub-collection S of S of size at most k such that Sj S S j = U. Given an instance (G = (V, E), k) of the Dominating Set problem, we can formulate it as an instance of the Set Cover problem by taking U = V and S = {S v = N[v] v V }. It is easy to verify that G has a dominating set of size k if and only if (U, S) has a set cover of size at most k. Hence the parameterized version of the Set Cover problem is W [2]-complete [8]. Here, we show that a special case of the Set Cover problem, that generalizes the Dominating Set problem for G 5 graphs to be fixed parameter tractable. More specifically, we show if the Set Cover instance (U, S) satisfies the property that for any pair of sets S i and S j in S, S i S j c, for a fixed constant c, then the problem is fixed parameter tractable. We call this variant of the Set Cover problem, where every pair of sets in the given family intersect in at most c elements, as Bounded Intersection Set Cover (BISC) problem. Observe that if the input graph G of the dominating set problem is a G 5 graph then the sets in its corresponding set cover instance satisfies a property that for any pair of sets S u and S v in S, S u S v 2. Theorem 8. The BISC problem is fixed parameter tractable. Proof. If there is a set S i S such that S i > ck then S i must be in every k-sized set cover. Otherwise, we need more than k sets to cover all the elements of U since every other set can cover at most c elements of S i. So this gives us a following simple reduction rule: Rule 1: Given a set cover instance, (U, S, k), if there exists S i S such that S i > ck then obtain a new reduced instance of set cover as following: U U S i. S {S S i S S}. If there are multiple copies of some set, then remove all but one copy of the same. k k 1 If k becomes 0 and U is non-empty then this is a no instance for the problem and we stop. We apply the Rule 1 until all the sets in S is of size at most ck, where k k. As k sets of size ck can only cover at most ck 2 ck 2 elements of U, if U > ck 2 then it is a no instance of the

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