The following is one of the important problems in MANETs.
WiFi interface is known to be a primary energy consumption in mobile devices. The “idle listening” consumes more energy compared with transmission or reception. It is estimated that about 60% of the energy is wasted in idle listening. The only solution for reducing idle listening is to implement a sleep schedule. Thus, intuitively whenever the nodes in S are communicating, u can go to sleep mode to conserve its battery power from idle listening.
MANET can be modeled as a simple undirected graph. An edge between the mobile hosts u and v indicates that both u and v are within their radio ranges. To simplify our discussion, we assume that if u is within the transmission range of v, then v will be in the transmission range of u (i.e., the relation between the nodes are symmetric. This may not be true in general due to different factors like the transmission power of a node, geographical conditions, etc.; thus, the resulting graph is a directed graph). Thus, the resulting graph is undirected.
We formally define the g-convexity and various parameters that are defined through the g-convexity.
From the above definition, it easily follows that a singleton set, vertex pair of an edge, and the whole vertex set are g-convex sets of G. We call them as trivial g-convex sets. Moreover, if S is a clique (S induces a complete subgraph of G), then S is a g-convex set of G.
2.1. The G-Centroid Location Problem for Arbirtary Connected Graphs
The g-centroid location has several practical applications. One such application domain is MANET. It also has an application in measuring dissimilarities and information retrieval [
5]. Due to its practical applications, it is thus necessary to devise an efficient algorithm to locate the g-centroid for an arbitrary graph. In this subsection, we outline the
-hardness of the g-centroid location algorithm. For the detailed proof, the readers may refer to our original paper [
2].
If the context is clear, in what follows, by the term graph we always mean a connected graph.
The following proposition specifies the structure of a g-centroid and its convexity.
Proposition 1. For any connected graph G, is a g-convex set of G and is connected.
Based on this proposition, we have the following results.
Proposition 2. For a connected graph G, lies in a block of G.
We now define the k-th neighborhood of a vertex u.
Let G = (V,E) be a connected graph and . The k-th neighborhood of u, denoted by consists of all vertices in G that are at a distance k from u, i.e., = {: }.
The next result is obvious from the definition of the k-th neighborhood of a vertex and its maximal g-convex set realizing its weight.
Proposition 3. Let G be a connected graph and u be a vertex of G. If , then .
The following corollary is immediate from Proposition 3.
Corollary 1. Let be a connected graph. For every vertex u in G, is either empty or induces a complete subgraph of G.
Note that for a vertex u of G, may be empty. As a nice open problem it will be interesting to classify all graphs for which , for every vertex u and any arbitrary . One such class is a tree.
We now outline the -hardness of the g-centroid location algorithm. The proof is by polynomially reducing the “clique decision” problem to the g-centroid location problem. However, we could not establish the membership of the g-centroid location problem in -class to establish the -completeness. g-convexity is closely related to the clique incident pattern of the graph.
We recall the definition of the “clique decision problem”:
Given a connected graph G and an integer r with , does G has a clique of size r?
The clique decision problem is one of the classical
-complete problem in graph theory. Several graph-theoretic and optimization problems were proved to be
-complete or
-hard by reducing to the clique decision problem [
7].
Definition 3. For a given connected graph G and an integer r with , we construct the graph from a copy of G, (the complete graph on vertices) and three new vertices a, b, and c as follows.
=
The edge set of consists of all the edges of G, , and the following new edges.
- -
Join a and b to all the vertices of G.
- -
Join c to a and b and to some arbitrary vertex d of .
For a given graph
G, the graph
will look as in
Figure 1. Furthermore, it is easy to see that for a given graph
G and a
r,
,
can be constructed in polynomial time. We explain this polynomial-time construction now:
We may assume that the graph is stored as an “adjacency matrix”. To obtain , we need to add vertices that corresponds to and the three new vertices a, b, and c. This is done by adding rows and columns to the “adjacency matrix”. Creating adjacency entries that represents takes time. Joining a to all the vertices of G is obtained by setting 1 for all columns that correspond to the vertices of G. This can be done in linear time. Similarly, joining b to all vertices of G can be done in linear time. Joining c to a and b and to some arbitrary vertex d of can be done in a constant time. Thus, the entire construction of from G takes polynomial time.
For an arbitrary connected graph G, we now analyze the structure of the g-centroid of for various values of r.
In what follows in this section, if the context is clear, we assume that the graph under consideration is .
The following facts can easily be established.
Proposition 4. Let G be a connected graph and be defined as in Definition 3. Let x be a vertex in the copy of of , then .
The proof follows from the fact that is complete and therefore .
Proposition 5. Let G be a connected graph and be defined as in Definition 3. .
The next proposition specifies the weight of the two vertices a and b based on the maximum clique size of G.
Proposition 6. Let G be a connected graph and be defined as in Definition 3. , where is the maximum clique size of G.
The following proposition determines the weight of the vertex c based on the chosen r and the maximum clique size of G.
Proposition 7. Let G be a connected graph and be defined as in Definition 3. = max {, }, where .
We now determine the weight of every vertex . The weight of these vertices depends on the maximum clique incident pattern of G and the chosen r.
Proposition 8. Let G be a connected graph and be defined as in Definition 3. Let be the maximum cliques of G, , and . Then, the following hold.
If , then for every , = max {,}.
If , then for every , = max {w,}. For every , = max {,}.
From the above propositions, for a given connected graph G and an integer r with , we can find the weight of every vertex of .
The following proposition analyze the structure of for various values of r.
Proposition 9. Let G be an arbitrary graph with the maximum clique size = w. Let r be an integer such that . Let be defined as in Definition 3. Then, the g-centroid of , = or M, depending upon whether the intersection of all the maximum cliques of G denoted by M is empty or not.
The following two propositions relate the chosen r and the maximum clique size of G.
Proposition 10. Let G be a connected graph with the maximum clique size = w and be defined as in Definition 3. Let . Then, = {c}.
Proposition 11. Let G be a connected graph with the maximum clique size = w and be defined as in Definition 3. Let . Then, = , irrespective of whether M is empty or not.
Combining all the results for , we have the following theorem.
Theorem 1. Let G be any connected graph and r be an integer such that . Let be defined as in Definition 3. Let be the maximum clique size of G. If , then = M or depending upon whether the intersection of all the maximum cliques of G denoted by M is non-empty or not. If , = , , = {c} irrespective of whether M is empty or not.
Based on Theorem 1, we can address the clique decision problem in polynomial time.
If , then the g-centroid of is either or M depending upon whether the intersection of all the maximum clique of G is empty of not. For , the g-centroid = . For , . Thus, G has a clique of size k if and only if or M.