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On a greedy approach for genome scaffolding
Algorithms for Molecular Biology volume 17, Article number: 16 (2022)
Abstract
Background
Scaffolding is a bioinformatics problem aimed at completing the contig assembly process by determining the relative position and orientation of these contigs. It can be seen as a paths and cycles cover problem of a particular graph called the “scaffold graph”.
Results
We provide some NPhardness and inapproximability results on this problem. We also adapt a greedy approximation algorithm on complete graphs so that it works on a special class aiming to be close to real instances. The described algorithm is the first polynomialtime approximation algorithm designed for this problem on noncomplete graphs.
Conclusion
Tests on a set of simulated instances show that our algorithm provides better results than the version on complete graphs.
Background
Motivation
In this paper, we focus on a bioinformatic problem occurring in the production of genomes. Genomes are usually obtained by sequencing. Sequencing produces an important amount of small sequences of nucleotides called reads. Herein, the lengths range from hundreds to tens of thousands of characters, depending on the sequencing technology. As a rule of thumb, shorter reads, produced for example by second generation sequencing (Illumina) have a higher quality (contain less readerrors) than long reads produced by third generation sequencing technologies (PacBio or Oxford Nanopore) [1]. The assembly process exploits overlaps between reads to reconstruct the targeted sequence. However, this is complicated by repeated parts in realworld genomes. Assembly algorithms cannot uniquely infer the original genome if it contains such repeated regions (the longer the repeated region with respect to the read length, the harder it is to infer the original genome). To avoid misassembly, such algorithms reconstruct only parts of the genome which is then returned as set of “contiguous regions” (or contigs). A thus fragmented genome is not ideal for further processing, and one would like to have as few contigs as possible while avoiding misassembly. A way to approach this are hybrid strategies using both long and short reads [2]. However, many genomes comprising current databases have been assembled before the development of third generation sequencing, preventing such hybrid strategies. One way to reduce the fragmentation of genomes in these databases while avoiding costly resequencing, is the exploitation of “metainformation” about the available reads.
Genome scaffolding
In second generation sequencing, short reads come in pairs, indicating that a fragment of the DNA molecule exists whose ends correspond to the two reads of a pair. In particular, the total length of said fragment is known approximately. This pairing information can be used to infer the order (and orientation) of the given contigs on the chromosome, thus completing the genome (modulo possible gaps between the contigs). The mathematical problem modeling this inference, called scaffolding, is made complicated by possible inconsistencies in the pairing information. See [3] for a recent overview of models, variants, and methods in this context.
The problem we study here is an optimization problem in a special graph called scaffold graph. The present formulation use both pairing information and some genomic structural constraints, like a fixed number of linear and circular chromosomes. In [4], we presented preliminary results about the complexity of this problem and a first polynomialtime approximation on complete graphs. Those results were extended and completed by another polynomialtime algorithm [5] and by a randomized approach [6]. We also explored exact algorithms [7], and studied some sparse special cases of scaffold graphs [8]. The contribution of the present paper is a continuation of published works [9, 10], where special classes of graphs have been studied (from sparse to very dense). Since real instances are usually sparse but contain some dense regions, due to abundance of repeats [11], we are interested in graphs built from cliques that are separated by bridges (i.e. edges whose removal disconnects the graph). The main contribution is the extension of the approximation algorithm on complete graphs of Chateau and Giroudeau [5] to a particular class called “connected cluster graph ”. Ultimately, the objective is to adapt the algorithm to sparse classes of graphs. To keep the approximation algorithm in polynomial time, one condition is that the decision problem of the scaffolding must be solvable in polynomial time. We propose a negative result, (i.e. it is \(\mathcal{NP}\mathcal{}\)complete) for a particular sparse graph class. Finally, since the presented approximation has a polynomial approximation ratio in some particular cases, we show that the scaffolding problem can not be approximed with a ratio better than a polynomial function in such cases.
Organization of the paper
The next section is devoted to notations and the description of the scaffold problem. In "Computational hardness" section, we show a \(\mathcal{NP}\mathcal{}\)hardness results for sparse scaffolding graphs. In "Nonapproximability" section, we address inapproximability. "Feasibility function for connected cluster graphs" section is devoted to a greedy algorithm for a special class of graph called connected cluster graph. Finally, we provide experimental results for the greedy algorithm.
Notation and problem description
Graph definitions
For a graph G, we denote by V(G) and E(G) the set of vertices and edges of G, respectively. Let u be a vertex of G, the degree d(u) of u is the number of edges incident with u. The girth g(G) of G is the length of the smallest cycle of G. A graph is bipartite if its vertices can be partitioned into two sets of nonadjacent vertices. A graph is planar if it can be drawn in the twodimensional plane without crossing edges.
A matching \(M^*\subseteq E(G)\) of G is a set of nonadjacent edges. \(M^*\) is called perfect if it touches all vertices of G. For a vertex u, we let \(M^*(u)\) denote the unique vertex v (if it exists) such that \(uv \in M^*\). In a scaffold graph, vertices represent extremities of contigs. Given a matching \(M^*\), the matching edges represent contigs and edges outside the matching represent possible contiguity relationship between contigs. The confidence that two contigs (more precisely, contigextremities) occur consecutively in the genomic sequence is represented by a weight on edges outside the matching. An alternating path (resp. alternating cycle) is a path (resp. cycle) such that its edges alternatingly belong to \(M^*\) or not. The extremal edges of an alternating path must be in \(M^*\).
A clique of G is a set of vertices such that all vertices are adjacent. A bridge (resp. cut vertex) of G is an edge (resp. vertex) such that its deletion increases by one the number of connected components of G. In "Feasibility function for connected cluster graphs" section, we study a particular class of graph called connected cluster graph, defined as follows.
Definition 1
A connected cluster graph G is a graph that admits a decomposition of its edges \(E(G)=E'\cup B\) such that the subgraph induced by \(E'\) is a disjoint union of cliques and each edge \(e \in B\) is a bridge of G.
An example of a connected cluster graph is given in Fig. 1.
Scaffolding problem
A scaffold graph \((G^*,M^*,\omega )\) is a simple, loopless graph \(G^*\) with a perfect matching \(M^*\) and a weight function \(\omega\) on the nonmatching edges. The matching \(M^*\) represents the set of contigs and for an edge uv, \(\omega (uv)\) indicates the confidence that the contig extremity v follows the contig extremity u in the genomic sequence.^{Footnote 1} The alternating girth of a scaffold graph denoted by \(g^*(G^*)\) is the number of matching edges in the smallest alternating cycle of \((G^*,M^*,\omega )\). In this paper, we study a decision and optimization version of scaffolding, defined as follows.
The two integers \(\sigma _p\) and \(\sigma _c\) are used to model restrictions on the sought genomic structure by representing the number of linear and circular chromosomes, respectively.
Let S be a collection of p alternating paths and c alternating cycles. We call the number \(p+c\) the cardinality of S and, we let \(\sigma _p(S):=p\) and \(\sigma _c(S):=c\).
Greedy algorithm
The main contribution of this paper is an extension of a known polynomialtime 3approximation [5] to connected cluster graphs. Whereas the original algorithm was developed to work in complete graphs, it can be adapted for the general case, as shown in Algorithm 1. The idea of this greedy algorithm is to consider each nonmatching edge in decreasing order of weight and add it into a partial solution, if possible. The key instruction is the feasibility function: given a partial solution S and an edge e, this function indicates whether \(S\cup e\) can still be extended into a collection of \(\sigma _c\) alternating cycles and \(\sigma _p\) alternating paths in \((G^*,M^*)\).
Proposition 1
Let f be a feasibility function with time complexity \(\mathcal {O}(t)\). Algorithm 1 gives an approximate solution for max scaffolding (if it exists) in \(\mathcal {O}(E(G^*) \cdot (t+\log E(G^*)))\).
The solution S given in the input of the feasibility function is called initiating solution. In general, since scaffolding is \(\mathcal{NP}\mathcal{}\)complete, feasibility cannot be decided in polynomialtime, even if \(S=\varnothing\) (unless \(\mathcal {P}\) =\(\mathcal{NP}\mathcal{}\)). Thus, we focus on restricted classes of graphs. In [5], a constanttime feasibility function was developed for complete graphs, leading to the following result.
Theorem 1
([5]) In complete graphs, Algorithm 1 gives a solution with an approximation factor of 3.
In "Feasibility function for connected cluster graphs" section, we develop a feasibility function for connected cluster graphs and show that Algorithm 1 gives a 5approximate solution in this case. Notice that, on graph classes containing the \(2\times k\) grids, the worstcase approximation factor of the greedy algorithm cannot be better than polynomial, even if a polynomialtime feasibility function exists (see Fig. 2).
We conclude this section with a note on realworld instances, which are too sparse to fall into our considered class. However, we can transform them by adding some nonmatching edges with weight zero. This technique was used to run the feasibility function for complete graphs on simulated instances [5] and the computed solution was close to the optimal. One of the reasons we develop a feasibility function for connected cluster graphs is that we conjecture that using a feasibility function for a graph class that is closer to the original instance (edgedeletion distance from the class) provides better approximation in practice, even though the theoretical approximation factor of the algorithm becomes worse. We test this hypothesis in "Experimental results" section.
Computational hardness
Like said in the previous section, when using the greedy algorithm on a real instance, we must complete the original instance by adding nonmatching edges with weight zero. To minimize the number of added edges, the solution is to adapt the greedy algorithm to a sparse class of graphs. In order to do that, scaffolding must be solvable in polynomial time in this particular class since otherwise, the feasibility function can not be run in polynomial time. In this section, we show that scaffolding is \(\mathcal{NP}\mathcal{}\)hard for the particular class of graphs where \(M^*=2\sigma _c + \sigma _p\). That is, we show that the greedy algorithm can not be executed in polynomial time in this special case. In such instance, any feasible solution S contains only alternating paths of length one and alternating cycles of length four (i.e. the smallest possible elements). While scaffolding is polynomial in this case [5], a natural extension would be to consider slightly longer alternating paths and alternating cycles. Unfortunately however, it turns out that deciding whether \((G^*,M^*)\) contains a collection with alternating paths of length one and alternating cycles of length six is already \(\mathcal{NP}\mathcal{}\)complete. In order to show this, we focus on the value of the alternating girth of the scaffold graph. Indeed, in a solution of scaffolding with \(g^*(G^*) \cdot \sigma _c + \sigma _p\) edges, each alternating path consists of exactly one matching edge and each alternating cycle is an alternating girth. We show that finding such a solution is \(\mathcal{NP}\mathcal{}\)complete, even if \(g^*(G^*)=3\), by reducing independent set to it.
IS is \(\mathcal{NP}\mathcal{}\)complete in general graphs. In order to build our reduction, we need G to be subcubic and trianglefree (i.e. \(\Delta (G)\le 3\) and \(g(G) > 3\)). Note that Lozin et Milanič [12] showed that independent set remains \(\mathcal{NP}\mathcal{}\)complete in \(\mathcal {F}\)free planar subcubic graphs if \(\mathcal {F}\) does not contain a tree with exactly three leaves. By choosing \(\mathcal {F}:=\{C_3\}\) (where \(C_3\) is the cycle on three vertices), we obtain the desired \(\mathcal{NP}\mathcal{}\)completeness. Our reduction uses the following construction.
Construction 1
(see Fig. 3) Given a subcubic, trianglefree graph G, construct a scaffold graph \((G^*,M^*,\omega )\) as follows:

for each edge \(e_i \in E(G)\), construct a matching edge \(u_i\overline{u}_i\), and

for each vertex \(v_t \in V(G)\), introduce the matching edges \(\{u^j_t\overline{u}^j_t \mid j \le 3deg(v_t)\}=:E_t\) and construct an alternating 6cycle \(C_t\) on the vertices \(E_t \cup \{u_i\overline{u}_i \mid v_t \in e_i\}\) such that no two u (or \(\overline{u}\)) vertices are adjacent.
The alternating cycles \(C_i\) are called vertexcycles. A bipartition is given by the u and \(\overline{u}\)vertices. Note that, if G is planar, it is also possible to construct a planar graph (which may no longer be bipartite). To show hardness of scaffolding when \(M^*= g^*(G^*)\cdot \sigma _c + \sigma _p\), we use the following properties of graphs resulting from Construction 1.
Lemma 1
Let G be a subcubic trianglefree graph and let \((G^*,M^*,\omega )\) be its scaffold graph produced by Construction 1. Let S be a collection of \(\sigma _c = k\) alternating cycles and \(\sigma _p= M^*3k\) alternating paths. Then,

(a)
\(g^*(G^*)=3\),

(b)
every alternating cycle in S is a vertexcycle, and

(c)
let \(C_t\) and \(C_{t'}\) be vertexcycles in S, the vertices \(v_t\) and \(v_{t'}\) are not adjacent in G.
Proof

(a)
By construction, each vertexcycle contains exactly three matching edges and, thus, \(g^*(G^*) \le 3\). Suppose there is an alternating cycle containing exactly two matching edges e and \(e'\). Let \(C_t\) be a vertexcycle containing e. Since \(C_t\) has length six, there is another vertexcycle \(C_{t'}\ne C_t\) that contains \(e'\). Indeed, e and \(e'\) are both in \(C_t\) and \(C_{t'}\) since, otherwise, their extremities cannot be adjacent. By construction, there are two edges \(e_i\) and \(e_j\) in G that are incident to both \(v_t\) and \(v_{t'}\), contradicting G being simple. Hence, there is no alternating cycle with two matching edges and \(g^*(G^*)=3\).

(b)
Let C be an alternating cycle in S. By Lemma 1(a), \(M^* = g^*(G^*)\cdot \sigma _c + \sigma _p\), implying that C has length six. Let \(u_i\overline{u}_i\) be a matching edge of C. If there is a matching edge \(v^1_t\overline{v}^1_t \in C\) then, by construction, the third matching edge of C is either \(v_t^2\overline{v}_t^2\) (if \(deg(v_t)=1\)) or \(u_j\overline{u}_j\) (where \(v_t \in e_j\) in G). Thus, C is the vertexcycle \(C_t\). Suppose there is no matching edge \(v^1_t\overline{v}^1_t\) in C. For any pair of matching edges \((u_k\overline{u}_k,u_{k'}\overline{u}_{k'})\) of C, \(e_k\) and \(e_{k'}\) are incident to a same vertex in G. Let \(u_i\overline{u_i},u_j\overline{u_j}\) and \(u_k\overline{u_k}\) be the three matching edges of C. Since G is trianglefree, \(e_i, e_j\) and \(e_k\) are adjacent in G, hence, C is a vertexcycle.

(c)
Let \(e_i = v_tv_{t'}\in E(G)\). The matching edge \(u_i\overline{u}_i\) is in \(C_t\) and \(C_{t'}\) and, thus, S cannot contain both \(C_t\) and \(C_{t'}\).
□
In the proof of correctness, we simulate vertices of the independent set with vertexcycles. If a solution S contains two vertex cycles \(C_i\) and \(C_j\), then \(v_i\) and \(v_j\) are not adjacent in G. Hence, if a solution S contains k vertexcycles, then there is an independent set of k vertices in G.
Theorem 2
scaffolding is \(\mathcal{NP}\mathcal{}\)complete, even in bipartite (or planar) subcubic scaffold graphs \((G^*,M^*,\omega )\) were \(M^* = g^*(G^*) \cdot \sigma _c + \sigma _p\) and \(g^*=3\).
Proof
Since, clearly, scaffolding is in \(\mathcal{NP}\mathcal{}\), it remains to show that Construction 1 is a reduction, that is, G has an independent set of size k if and only if there is a collection of k alternating cycles and \(M^*3k\) alternating paths in \((G^*,M^*)\).
“\(\Rightarrow\)”: Let I be an independent set of size k in G. We build a solution of scaffolding as follows. For each vertex \(v_t \in I\), we construct the vertexcycle \(C_t\) in S. For each remaining matching edge in \(M^*\setminus \bigcup _{v_t\in I}C_t\), we construct an alternating path of length one. We obtain a solution S as thought.
“\(\Leftarrow\)”: Let S be a solution in \((G^*,M^*)\) containing k alternating cycles and \(E(G)k\) alternating paths and let \(I := \{v_t  C_t \in S\}\). By Lemma 1(b), any alternating cycle of S is a vertexcycle in \((G^*,M^*)\) and, thus, \(I=k\). Moreover, by Lemma 1(c), I is independent in G.□
Note that Theorem 2 can be generalized to \(g^*(G^*)>3\) by modifying Construction 1 as follows. First, we build our construction from a graph G with \(g(G)>\ell \ge 3\). IS remains \(\mathcal{NP}\mathcal{}\)complete in such graphs by the result of Lozin and Milanič: it suffices to take \(F=\{C_{i} \mid i \le \ell \}\), where \(C_i\) is the cycle of order i. Then, we increase the length of every vertexcycle by taking \(E_t = \{u^j_t\overline{u}^j_t \mid j\le 3+\ell  deg(v_t)\}\) for each \(v_t \in V(G)\). By making these modifications, we construct a scaffold graph with \(g^*(G^*) = \ell\) and we preserve properties Lemma 1(b) and Lemma 1(c). This leads to the following result.
Corollary 1
scaffolding is \(\mathcal{NP}\mathcal{}\)complete even in bipartite (or planar) subcubic scaffold graphs \((G^*,M^*)\) were \(M^* = g^*(G^*) \cdot \sigma _c + \sigma _p\), for all \(g^*(G^*)\ge 3\).
Nonapproximability
In this section, we discuss the hardness of approximating max scaffolding. Notice that, since scaffolding is \(\mathcal{NP}\mathcal{}\)complete, there is no polynomialtime approximation algorithm for max scaffolding (unless \(\mathcal {P} =\mathcal{NP}\mathcal{}\)). However, this argument does not hold for graph classes where scaffolding is in \(\mathcal {P}\) (i.e. classes for which the feasibility function (and, thus, the greedy algorithm) runs in polynomial time).
We show that, in this case, max scaffolding is still PolyAPXhard, that is, it is not possible to approximate max scaffolding within a factor better than a polynomial function in \(V(G^*)+E(G^*)\) (unless \(\mathcal {P} =\mathcal{NP}\mathcal{}\)). Recall that Fig. 2 already shows that the greedy algorithm can not approximate max scaffolding with a ratio better than a polynomial function. The inapproximability result presented in this section shows that it is the case for any polynomialtime algorithm. In the following, we construct an Sreduction (see [13]) from the optimization version of independent set.
Construction 2
(see Fig. 4) Let G be a graph. Then, construct the following scaffold graph \((G^*,M^*,\omega )\):

For each \(e_i= v_tv_q\in E(G)\), construct a clique \(\{u_i^t,\overline{u}_i^t,u^q_i,\overline{u}_i^q,e_i,\overline{e}_i\}\) with \(u_i^t\overline{u}_i^t,u_i^q\overline{u}_i^q,e_i\overline{e}_i \in M^*\).

For each \(v_t\in V(G)\), construct a cycle \((v^t_1,\overline{v}^t_1,\overline{v}^t_2,v^t_2)\) with \(v^t_1\overline{v}^t_1,v^t_2\overline{v}^t_2 \in M^*\).

Let \(v_t \in V(G)\) and let \(\mathcal {A}_t\) be a list of all edges incident with \(v_t\) in G. Construct an alternating cycle containing all vertices in \(\{v^t_1,\overline{v}^t_1,v^t_2,\overline{v}^t_2\} \cup \{ u^t_i,\overline{u}^t_i \mid \forall e_i \in \mathcal {A}_t\}\) as follows:

For all \(k < d(v_t)\), let \(e_i\) and \(e_j\) be the \(k^\text {th}\) and \(k+1^\text {st}\) edges of \(\mathcal {A}_t\), respectively, and add a nonmatching edge between \(\overline{u}^t_i\) and \(u^t_j\).

Let \(e_i\) and \(e_j\) be the first and last edges of \(\mathcal {A}_t\), respectively, and add the nonmatching edges \(v^t_1u^t_i\) and \(v^t_2\overline{u}^t_j\).


Each nonmatching edge has weight zero, except the edges \(v^t_2\overline{e}_j\) which have weight one.
Let \(v_t\in V(G)\). The cycle on \(\{v^t_1,\overline{v}^t_1,v^t_2,\overline{v}^t_2\} \cup \{u^t_i,\overline{u}^t_i \mid \exists q\; e_i = v_tv_q \in E(G)\}\) is called the long vertexcycle of \(v_t\) and is denoted by \(C(v_t)\). Note that a long vertexcycle has weight one. Now consider the following properties.
Lemma 2
Let G be a graph and let \((G^*,M^*,\omega )\) be the scaffold graph produced by Construction 2. Let S be a collection of \(V(G)+E(G)\) alternating cycles in \((G^*,M^*,\omega )\).

(a)
Every nonzeroweight alternating cycle C of S is a long vertexcycle.

(b)
Let \(C(v_t)\) and \(C(v_q)\) be two long vertexcycles of S. Then, \(v_tv_q\notin E(G)\).
Proof
Note that it is always possible to build a collection of \(V(G)+E(G)\) (weight0) alternating cycles in \((G^*,M^*,\omega )\) by constructing the alternating cycle \(\{u^t_i, \overline{u}^t_i,u^q_i,\overline{u}^q_i,e_i,\overline{e}_i\}\) for each edge \(e_i=v_tv_q\) of G and the alternating cycle \(\{v^t_1,\overline{v}^t_{1},v^t_2,\overline{v}^t_2\}\) for each vertex \(v_t\in V(G)\).□
Claim 1
Let \(v_t\in V(G)\) and \(e_i\in E(G)\). Then, no alternating cycle of S contains both \(e_i\overline{e}_i\) and \(v^t_1\overline{v}^t_1\).
Proof
Towards a contradiction, assume that there is such an alternating cycle C. By pidgeonhole principle, one of the \(V(G)+E(G)\) alternating cycles in S, say \(C'\), avoids both \(e_i\overline{e}_i\) and \(v^t_1\overline{v}^t_1\) for all \(i,t\in \mathbb {N}\). Let \(u^t_i\overline{u}^t_i\) be a matching edge of \(C'\) for some \(e_i=v_tv_q\). Then, \(C'\) cannot contain \(u^q_i\overline{u}^q_i\) as, otherwise, \(e_i\overline{e}_i\) cannot be part of an alteranting cycle in S, implying that S is not a solution. Thus, each matching edge of \(C'\) is on the long vertexcycle \(C(v_t)\). Since the graph induced by the vertices of \(C(v_t) \setminus v^t_1\overline{v}^t_1\) is a path, it is not possible to construct \(C'\). Hence, we conclude that C does not exist.□
(a): Let C be a nonzeroweight alternating cycle of S and assume towards a contradiction that C is not a long vertexcycle. Since C contains a nonzeroweight edge \(v^t_2\overline{u}^1_i\), the matching edge \(v^t_2\overline{v}^t_2\) is in C. As C is not a long vertexcycle, there is some \(e_i=v_tv_q\) such that C contains both \(u^t_i\overline{u}^t_i\) and \(u^q_i\overline{u}^q_i\). Thus, either the matching edge \(e_i\overline{e}_i\) is in C, contradicting Claim 1, or \(e_i\overline{e}_i\) consists of a singleedge alternating path of S, contradicting our choice of S.
(b): Towards a contradiction, assume that S contains \(C(v_t)\) and \(C(v_q)\) such that \(e_i=v_tv_q\in E(G)\). Then, the matching edge \(e_i\overline{e}_i\) is a singleedge alternating path of S, contradicting our choice of S.
We now show the PolyAPXhardness of max scaffolding, even for graph classes for which \(\textsc {Scaffolding} \in \mathcal {P}\). Reusing the same idea of Theorem 2, we simulate the vertices of the independent set with long vertexcycles. If a solution S of max scaffolding has weight k, then S contains k long vertexcycles and, since their related vertices cannot be adjacent, we can construct an independent set with k vertices in G.
Theorem 3
max scaffolding is PolyAPXhard, even for graph classes for which \(\textsc {Scaffolding} \in \mathcal {P}\).
Proof
Let G be an instance of independent set and let \((G^*,M^*,\omega )\) be the scaffold graph produced by Construction 2. Let \(\mathcal {S}\) be the set of all collections of \(\sigma _p=0\) alternating paths and \(\sigma _c = V(G)+E(G)\) alternating cycles in \((G^*,M^*,\omega )\).
Recall that independent set is PolyAPXcomplete for general graphs [14]. We show that G has a sizek independent set if and only if \(\mathcal {S}\) contains a solution S of score k.
“\(\Rightarrow\)”: Let I be an independent set of size k in G. We construct a solution \(S\in \mathcal {S}\) as follows.
First, for each \(v_t\in I\), construct the alternating cycle \(C(v_t)\) in S. Second, for each \(v_t\in V(G)\setminus I\), construct the alternating cycle \((v^t_1,\overline{v}^t_{1},\overline{v}^t_2,v^t_2)\) in S.
Third, for each edge \(e_i=v_tv_q\) not incident with a vertex in I, construct the alternating cycle \((u^t_i, \overline{u}^t_i,\overline{u}^q_i,u^q_i,e_i,\overline{e}_i)\) in S.
Fourth, for each edge \(e_i=v_tv_q\) with \(v_t\in I\), (the matching edge \(u^t_i\overline{u}^t_i\) is in \(C(v_t)\) which is already in S), construct the alternating cycle \((u^q_i,\overline{u}^{q}_i,e_i,\overline{e}_i)\).
Since each long vertexcycle has weight one, we obtain a solution S with \(\omega (S)= k\).
“\(\Leftarrow\)”: Let \(S\in \mathcal {S}\) with \(\omega (S) = k\). We construct an independent set I by taking all vertices whose long vertexcycle is in S, that is, \(I := \{ v_t \mid C(v_t) \in S\}\). Since each long vertexcycle has weight one, Lemma 2a implies that S contains k long vertexcycles. Thus, \(I = k\). Further, by Lemma 2b, I is independent.
Let f be the function corresponding to Construction 2 and let g be a function that computes an independent set in G from a solution in f(G), as described above. Suppose that there is a polynomialtime algorithm A with approximation factor \(\rho\) for max scaffolding. The approximation factor of \(g \circ A \circ f\) is equal to \(\rho\), thus Construction 2 constitutes an Sreduction. Nonapproximability results of independent set transfer to max scaffolding.□
Feasibility function for connected cluster graphs
In this section, we present a feasibility function for connected cluster graphs using dynamic programming. For simplicity, we consider in the following scaffold graphs \((G^*,M^*,\omega )\) such that \(G^*\) is a connected cluster graph and no matching edge is a bridge. The case were a bridge can be a matching edge is included in the feasibility function for block graph that (see "Experimental results" section).
Definitions
Notice that the structure of a connected cluster graph is close to a tree (that is, collapsing each clique of \(G^*\) into a single vertex leads to a tree), so we will use a similar vocabulary: a rooted connected cluster graph is a connected cluster graph where a clique r is designated as a root. Then, the following notation applies: the parent of a clique x is the clique connected to x on the unique xrpath. A child of a clique c is clique of which c is the parent. Any clique without children is called a leaf. A vertex v of a clique c is a door of c if v is adjacent to a vertex u in a child of c. In that case, for simplicity, we say that the clique containing u is a child of v. The upper door of a clique \(c\ne r\) is the unique vertex v that is adjacent to a vertex of the parent of c. Let c be a clique of \(G^*\) and let S be a partial solution in \(G^*\). Let \(S'\) be the intersection of S and c, an alternating element of c is either an alternating cycle of \(S'\) or an alternating path of \(S'\). Notice that an alternating path of S can be decomposed into several alternating elements if it belongs to several cliques. Let e be the alternating element containing the upper door of c. The subclique \(c'\) of c is the subgraph containing every vertex of c that does not belong to e. Formally, \(c'=G^*[V(c)\setminus V(e)]\). We use the treestructure to develop a bottomup algorithm, that is, we construct and assemble some partial solutions from the leaves to the root. We define some operations to combine this partial solutions.
Operations
Let \(G_1\) and \(G_2\) be two edgedisjoint subgraphs. We can build a solution in the graph induced by \(V(G_1) \cup V(G_2)\) from a solution in \(G_1\) and a solution in \(G_2\), using four operations.
Definition 2
Let \(G_1\) and \(G_2\) be edgedisjoint subgraphs of \(G^*\). Let \(S_1\) and \(S_2\) be solutions of \(G_1\) and \(G_2\), respectively. Let S be a solution of \(G^*[V(G_1) \cup V(G_2)]\). S is a composition of \(S_1\) and \(S_2\) if S can be obtained from \(S_1\cup S_2\) by at most one of the following operations:
 Merger::

merge an alternating path \((u_1,u_2,\ldots ,u_{2t})\) of \(S_1\) with an alternating path \((v_1,v_2,\ldots ,v_{2q})\) of \(S_2\) by adding the nonmatching edge \(u_{2t}v_1\).
 Closing::

close an alternating path \((u_1,u_2,\ldots ,u_{2t})\) of \(S_1\) and an alternating path \((v_1,v_2,\ldots , v_{2q})\) of \(S_2\) into an alternating cycle by adding the nonmatching edges \(u_{2t}v_1\) and \(v_{2q}u_1\).
 Absorption::

replace a nonmatching edge \(v_{2i}v_{2i+1}\) of an alteranting path in \(S_2\) with an alternating path \((u_1,u_2,\ldots ,u_{2t}\) of \(S_1\) by removing \(v_{2i}v_{2i+1}\) and adding the nonmatching edges \(v_{2i}u_1\) and \(u_{2t}v_{2i+1}\). We call \(v_{2i}v_{2i+1}\) absorbent.
Finally, if no operation is necessary to obtain S from \(S_1\cup S_2\), we say that S is obtained by juxtaposition.
Note that all presented operations add only edges of \(E(G^*) \setminus (E(G_1) \cup E(G_2))\). Note further that not all compositions of two solutions are guaranteed to exist for a pair \(S_1\) and \(S_2\). In the algorithm, we manipulate sets of solutions: we can create a new set of solutions from two sets of solution if all pairs of solutions of the two input sets are used in the resulting set.
Definition 3
Let \(G_1\) and \(G_2\) be two edgedisjoint subgraphs of \(G^*\) and let \(\mathcal {S}_1\) and \(\mathcal {S}_2\) be sets of solutions of subgraphs \(G_1\) and \(G_2\), respectively. Let op be one the operation described in Definition 2. Then, we call the set \(\mathcal {S}=\{op(S_1,S_2)\mid \forall S_1\in \mathcal {S}_1 \wedge \forall S_2\in \mathcal {S}_2\;\}\) the complete composition of \(\mathcal {S}_1\) and \(\mathcal {S}_2\).
To ensure the possibility of building a complete composition from two sets of solutions, it is useful to characterize a solution according to the operations we can perform on it.
Definition 4
Let G and \(G'\) be two edgedisjoint subgraphs of \(G^*\) and let S be a feasible solution of scaffolding for \((G,M^*,\omega )\).

1.
S is closeable if S contains an alternating path \((u_1,u_2\ldots ,u_{2t})\) and \(G'\) contains an alternating path \((v_1,v_2,\ldots ,v_{2q})\) such that \(u_{2t}v_1\) and \(v_{2q}u_1\) are edges of \(E(G^*\setminus M^*\).

2.
S is extensible by \(G'\) if S contains a vertex v such that v is an extremity of an alternating path and v has a neighbor in \(G'\) .

3.
S is frozen to \(G'\) if S is not extensible.

4.
S is absorbent to \(G'\) if S contains an alternating path \((u_1,u_2,\ldots ,u_{2t})\) and \(G'\) contains an alternating path with extremities v and w such that \(vu_{2i}, wu_{2i+1} \in E(G^*)\setminus M^*\) for some \(i<t\). Note that an absorbent solution can also be closeable, alternating or frozen.
When omitted, \(G'\) defaults to \(G^* V(G)\).
Note that all closeable solutions are also extensible. If a solution S is closeable by a subgraph \(G'\), then we can close an alternating path of S into an alternating cycle by adding some edges of \(G'\). If a solution S is extensible by a subgraph \(G'\), then we can add some edges of \(G'\) in an extremity of an alternating path of S without changing the cardinality of the solution. Finally, if a solution S is absorbent to a subgraph \(G'\), then we can replace an absorbent edge of S by a path of length three without changing the cardinality of S. An example of the different operations of Definition 4 is given in Fig. 5.
Semantics
Since the number of possible solutions can be exponential, we just store the possible cardinalities in the table entries, which is sufficient to answer the question of feasibility. Recall that, if \(X, Y \subseteq \mathbb {N}\) are two sets of integers, then the sum of X and Y is defined as \(X+Y = \{ x+y\mid x \in X, y \in Y \}\). Note that \(X + \varnothing = \varnothing\). In the following, we call an integer j eligible with respect to a set \(\mathcal {S}\) of solutions and an integer i if there is a solution \(S \in \mathcal {S}\) containing i alternating cycles and j alternating paths. Then, our dynamic programming table has the following semantics.
Definition 5
(Semantics) Let \(\mathcal {S}\) be a set of solutions and let \(i \in \mathbb {N}\).
A table entry \([\mathcal {S},i]\) is the set of all integers eligible with respect to \(\mathcal {S}\) and i. More formally, letting \(X_i = \{S \mid S \in \mathcal {S} \wedge \sigma _c(S) = i\}\), we define \([\mathcal {S},i] = \{\sigma _p(S) \mid S \in X_i\}\).
Let us highlight three particular values of \([\mathcal {S},i]\). For \(\mathcal {S}=\{\varnothing \}\), we have \([\{\varnothing \},0] = \{0\}\) and, for each \(i>0\), we have \([\{\varnothing \},i]=\varnothing\). For an alternating path p, we have \([\{p\},0] = \{1\}\) and \([\{p\},i]=\varnothing\) for each \(i>0\). Finally, for an alternating cycle c, we have \([\{c\},1] = \{0\}\) and \([\{c\},i]=\varnothing\) for each \(i\ne 1\). For brevity, we let \([\mathcal {S}]\) denote the vector \(([\mathcal {S},0],\dots ,[\mathcal {S},\sigma _c])\) and, for any operator \(\diamond\) and any sets \(\mathcal {S}_1\) and \(\mathcal {S}_2\) of solutions, we define \([\mathcal {S}_1] \diamond [\mathcal {S}_2]\) as componentwise \(\diamond\), that is, \([\mathcal {S}_1,i] \diamond [\mathcal {S}_2,i]\) for each \(i \in [0,\sigma _c]\).
Lemma 3
Let \(G_1\) and \(G_2\) be two vertexdisjoint subgraphs of \(G^*\) and let \(\mathcal {S}_1\) and \(\mathcal {S}_2\) be sets of solutions of \(G_1\) and \(G_2\), respectively. Let \(\mathcal {S}\) be a set of solutions of \(G^*[V(G_1) \cup V(G_2)]\) such that \(\mathcal {S}\) is a complete composition of \(\mathcal {S}_1\) and \(\mathcal {S}_2\).

1.
If \(\mathcal {S}\) is the set of solutions composed with a merger operation, then \([\mathcal {S},k] = \bigcup _{i+j=k}([\mathcal {S}_1,i] + [\mathcal {S}_2,j] + \{1\})\).

2.
If \(\mathcal {S}\) is the set of solutions composed with a closing operation, then \([\mathcal {S},k] = \bigcup _{i+j+1=k}([\mathcal {S}_1,i] + [\mathcal {S}_2,j] + \{2\})\).

3.
If \(\mathcal {S}\) is the set of solutions composed with an absorption operation, then \([\mathcal {S},k] = \bigcup _{i+j=k}([\mathcal {S}_1,i] + [\mathcal {S}_2,j] + \{1\})\).

4.
If \(\mathcal {S}\) is the set of solutions composed with a juxtaposition operation, then \([\mathcal {S},k] = \bigcup _{i+j=k}([\mathcal {S}_1,i] + [\mathcal {S}_2,j])\).
Proof
Let \(S \in \mathcal {S}\) and let \(S_1\) and \(S_2\) denote the solutions of \(\mathcal {S}_1\) and \(\mathcal {S}_2\), respectively, such that S is composed by \(S_1\) and \(S_2\). Then,

1
since \(S_1\) and \(S_2\) have a common alternating path in S, we have \(\sigma _p(S) = \sigma _p(S_1) + \sigma _p(S_2) 1\) and since no cycle is formed, \(\sigma _c(S) = \sigma _c(S_1) + \sigma _c(S_2)\). Thus, since \(\mathcal {S}\) is a complete composition of \(\mathcal {S}_1\) and \(\mathcal {S}_2\), we have \([\mathcal {S},k] = \bigcup _{i+j=k}([\mathcal {S}_1,i] + [\mathcal {S}_2,j] + \{1\})\).

2
since one path of \(S_1\) and one path of \(S_2\) are closed into a single alternating cycle, we have \(\sigma _p(S) = \sigma _p(S_1) + \sigma _p(S_2) 2\) and \(\sigma _c(S) = \sigma _c(S_1) + \sigma _c(S_2) + 1\). Thus, since \(\mathcal {S}\) is a complete composition of \(\mathcal {S}_1\) and \(\mathcal {S_2}\), we have \([\mathcal {S},k] = \bigcup _{i+j=k}([\mathcal {S}_1,i] + [\mathcal {S}_2,j] + \{2\})\).

3
since \(S_1\) has an alternating path that is “absorbed” into a connected component of \(S_2\), we have \(\sigma _p(S) = \sigma _p(S_1) + \sigma _p(S_2) 1\) and since no cycle is formed, \(\sigma _c(S) = \sigma _c(S_1) + \sigma _c(S_2)\). Thus, since \(\mathcal {S}\) is a complete composition of \(\mathcal {S}_1\) and \(\mathcal {S_2}\), we have \([\mathcal {S},k] = \bigcup _{i+j=k}([\mathcal {S}_1,i] + [\mathcal {S}_2,j] + \{1\})\).

4
since all paths and cycles of \(S_1\) and \(S_2\) are present in S, we have \(\sigma _p(S) = \sigma _p(S_1) + \sigma _p(S_2) 1\) and since no cycle is formed, \(\sigma _c(S) = \sigma _c(S_1) + \sigma _c(S_2)\). Thus, since \(\mathcal {S}\) is a complete composition of \(\mathcal {S}_1\) and \(\mathcal {S_2}\), we have \([\mathcal {S},k] = \bigcup _{i+j=k}([\mathcal {S}_1,i] + [\mathcal {S}_2,j])\).
\(\square\)
We use Lemma 3 to define the four functions juxtapose, \(merge_t\), absorb, and \(close_t\), which provide table entries for complete compositions “composed” with a juxtaposition, merge, absorption or closing operation, respectively. Although Lemma 3 is defined for two sets, we use a generalized version which can take as parameters more than two sets. The functions \(merge_t\) and \(close_t\) have a parameter t that indicates the number of paths merged or closed during the operation. For example, if we have three sets \(S_1\), \(S_2\), and \(S_3\) and if it is possible to construct a single alternating path in the resulting composition by taking one alternating path in each set, then we use the function \(merge_3(\{S_1\},\{S_2\},\{S_3\})\). Note that the parameter t can be different from the number of sets. In addition, it is sometimes possible to close a single alternating path into an alternating cycle and, in that case, the function \(close_1\) is used. The four functions are defined in Algorithm 2, Algorithm 3 and Algorithm 4. However, we must ensure that the associated operation is feasible before using one these functions.
In the algorithm, we traverse four different types of subgraphs defined as follows.

Let \(v \in V(G^*)\), let child(v) be the set of children of v in \(G^*\) (possibly empty). Then, \(G^*(v)\) denotes the subgrah of \(G^*\) that is induced by v and every branch linked to v. Formally, \(G^*(v) := G^*[\{v\} \cup \bigcup \limits _{x \in child(v)} V(G^*(x))]\).

Let e be an alternating element. Then, \(G^*(e)\) denotes the subgraph of \(G^*\) that is induced by e and all children of its vertices. Formally, \(G^*(e) = G^*[\bigcup \nolimits _{v \in e} V(G^*(v))]\).

Let c be a clique of \(G^*\) and let \(c'\) be the subclique of c. For all \(x \in \{c,c'\}\), the subgraph \(G^*(x)\) is the union of x and all children of x. Formally, \(G^*(x)=G^*[\bigcup \nolimits _{e \in M^*\cap ( {\begin{array}{c}x\\ 2\end{array}}) } V(G^*(e))]\).
For each traversed subgraph, we use four different sets of solutions distinguishing solutions according to their properties.
Definition 6
Let S be a partial solution of \(G^*\). Let x be a vertex, a partial path, a subclique or clique of \(G^*\) and let \(S'\) be a solution of the subgraph \(G^*(x)\). Then,

\(S \in \mathcal {C}(x) \Leftrightarrow\) \(S'\) is closeable and \(S \cap E(G^*(x)) \subseteq S'\).

\(S \in \mathcal {E}(x) \Leftrightarrow\) \(S \notin \mathcal {C}(x)\) and S is extensible and \(S \cap E(G^*(x)) \subseteq S'\).

\(S \in \mathcal {A}(x) \Leftrightarrow\) S is frozen and absorbent and \(S \cap E(G^*(x)) \subseteq S'\).

\(S \in \mathcal {F}(x) \Leftrightarrow\) \(S \notin \mathcal {A}(x)\) and S is frozen and \(S \cap E(G^*(x)) \subseteq S'\).
The algorithm
We now present a method to provide the feasibility function needed by Algorithm 1. In the next paragraphs, we describe the algorithms that calculate the table entries for the four types of subgraphs described above.
Vertex
Let \(v\in V(G^*)\). We show in this part how to compute the table entries for the sets \(\mathcal {F}(v)\) and \(\mathcal {E}(v)\). Note that, since the edge between \(G^*(v)\) and its parent is a bridge, any solution \(S'\) for \(G^*(v)\) can have at most one edge incident to v. Thus, the sets \(\mathcal {C}(v)\) and \(\mathcal {A}(v)\) are empty. If v is not incident to an edge of \(S \cap E(G^*(v))\), then we construct the table entries by successively merging the table entries of the children adjacent to v. For that, we use at each step an intermediate graph \(G_i\). Let \(V_i\) be the set of the first i children of v. \(G_i\) is the subgraph of \(G^*\) induced by v and all vertices in \(V_i\). If v is incident with an edge \(S \cap E(G^*(v))\), then any solution containing S is in \(\mathcal {E}(v)\). An example of solutions computed by Algorithm 5 is depicted in Fig. 6.
Lemma 4
For any vertex v, the values of the table entries provided by Algorithm 5 are correct for the set \(\mathcal {F}(v)\) and \(\mathcal {E}(v)\).
Proof
First, if there is no child linked to v, then \(G^*(v)\) consists of the single vertex v. In that case, the only solution for \(G^*(v)\) consists of zero alternating cycles and paths and this solution is frozen. Thus, the initial values given to \([\mathcal {F}(v)]\) and \([\mathcal {E}(v)]\) in the initialization step (i.e. lines 1 to 2) are correct. Assume that table entries returned by \(compute\_clique\) are correct. Let \(S'\) be a solution of \(G^*(v)\) such that \(S\cap E(G^*(v)) \subseteq S'\). We distinguish two cases.
 Case 1::

there is an edge \(uv \in S\cap E(G^*(v))\). Thus, \(S'\) is extensible and is composed by the merge of an extensible solution in \(G^*(c_u)\) with uv and the juxtaposition of any solution for each child \(c_{u'}\ne c_u\). Hence, lines 9 and 11 are correct.
 Case 2::

there is no edge \(uv \in S\cap E(G^*(v))\). Then, \(S'\) is frozen if and only if it does not contain an edge incident to v. As there is no edge uv in any child \(c_t\), \(S'\) is composed by juxtaposition of any solution for each child \(c_t\) and the assignment in line 13 is correct. If \(S'\) is extensible, then there is a unique child \(c_t\) of v such that an alternating path from \(S' \cap E(G^*(c_t))\) has been expanded to v and, therefore, the solution \(S' \cap E(G^*(c_t))\) is extensible. Thus, \(S'\) is composed by a merge of a extensible solution of a unique child and the juxtaposition of any solution in other children. Hence, line 14 is correct.
□
Alternating element
Let c be a clique of \(G^*\) and let e be an alternating element of c such that e does not contain the upper door of c. We show in this part how to compute the table entries for the sets \(\mathcal {C}(e)\), \(\mathcal {F}(e)\) and \(\mathcal {E}(e)\). If e is a uvpath, then the idea is to merge the computed table entries of u and v and juxtapose the frozen solutions of the inner vertices. If e is an alternating cycle, then there is no choice and the only solution containing S is frozen. An example of solutions computed by Algorithm 6 is depicted in Fig. 7.
Lemma 5
For any alternating element e, the values of the table entries provided by Algorithm 6 are correct for the sets \(\mathcal {C}(e)\), \(\mathcal {F}(e)\) and \(\mathcal {E}(e)\).
Note that the only possibility to obtain an absorbent solution of \(G^*(e)\) is when e is a path that is closed into an alternating cycle. However, if an absorption operation is done in the function \(compute\_subclique\), then the resulting solution can also be obtained by a closing operation with a solution in \(\mathcal {C}(e)\). Thus, our dynamic programming will not compute the value of \([\mathcal {A}(e)]\).
Proof
Suppose that the values of the table entries provided by the function \(compute\_vertex\) are correct. First note that, for each inner vertex \(v_t\) of e, the subsolutions of \(G^*(v_t)\) are necessarily frozen, then a solution of \(G^*(e)\) contains a juxtaposition of frozen solutions of the inner vertices of e. If e is an alternating cycle, then the only possible solution is obtained by the juxtaposition of frozen solutions of the inner vertices and the alternating cycle e. Thus, the assignment line 5 is correct. Suppose that e is a partial path. All possible values of the juxtaposition of the frozen solutions of the inner vertices are assigned in the table entry \([\mathcal {I}_e]\).

A solution \(S'\) of \(G^*(e)\) is closeable if the degree of the extremities of e are equal to one. Then, the subsolutions of \(S'\) in \(G^*(v_0)\) and \(G^*(v_k)\) are frozen. Thus, the assignment line 7 is correct.

A solution \(S'\) of \(G^*(e)\) is frozen if the degree of the extremities of e are equal to two. It is the case if (1) the subsolutions of \(S'\) in \(G^*(v_0)\) and \(G^*(v_k)\) are extensible or (2) the subsolutions of \(S'\) in \(G^*(v_0)\) and \(G^*(v_k)\) are frozen and e is closed into an alternating cycle. Thus, the assignment line 8 is correct.

A solution \(S'\) of \(G^*(e)\) is extensible and not closeable if and only if exactly one vertex in \(\{v_0,v_k\}\) has degree one. Then, exactly one subsolution of S in \(G^*(v_0)\) or \(G^*(v_k)\) is extensible. Thus, the assignment line 9 is correct.
□
Subclique
Let \(c'\) be a subclique of \(G^*\) containing k alternating elements. We show in this part how to compute the table entries for the sets \(\mathcal {C}, \mathcal {F}, \mathcal {A}\) and \(\mathcal {E}\). The idea is to construct the table entry by merging successively each table entry of the alternating elements of \(c'\). For that, we use at each step an intermediate graph \(G_t\) and two intermediate sets \(\mathcal {A}_+\) and \(\mathcal {E}_+\), defined as follows. Let \(L(c') = \{e_1,\dots ,e_k\}\) be a list of alternating elements of \(c'\), let \(t\le k\), let \(E_t=\{e_1,\dots ,e_t\}\), and let \(V_t =\bigcup _{e \in E_t} V(G^*(e))\). Let \(G_t\) be the subgraph of \(G^*\) induced by \(V_t\). At step t, a solution \(S'\) is in \(\mathcal {A}_+\) (resp. \(\mathcal {E}_+\)) if and only if (1) \(S'\) is a solution of \(G_t\), (2) \(S'\) contains a set \(C \ne \varnothing\) of closeable paths and (3) \(S \setminus C\) is not extensible (resp. extensible).
Lemma 6
For any subclique \(c'\), the value of the table entries provided by Algorithm 7 are correct for the sets \(\mathcal {C}(c'), \mathcal {F}(c'), \mathcal {A}(c')\) and \(\mathcal {E}(c')\).
Proof
Assume table entries returned by \(compute\_alternating\_element\) are correct. We show by induction that the values calculated in each step t are correct for the graph \(G_t\). First, \(G_0\) is the empty graph and the unique solution is that containing zero alternating cycles and paths and this solution is frozen. Thus, lines 1 to 3 are correct. Now, consider the alternating element \(e_t\) and suppose the previously computed values are correct (i.e. values stored in \(\mathcal {F}',\mathcal {A}',\mathcal {E}',\mathcal {A}_+'\) and \(\mathcal {E}_+'\)). Let \(S_1\) be a solution in \(G_{t1}\), let \(S_2\) be a solution in \(G^*(e_t)\) and let \(S'\) be a composition of \(S_1\) and \(S_2\). We have the following properties:

if \(S'\) is obtained by a juxtaposition, then \(S_1\) is in \(\mathcal {F}',\mathcal {A}',\mathcal {E}',\mathcal {A}_+'\) or \(\mathcal {E}_+'\) and \(S_2\) is in \(\mathcal {C}(e_t),\mathcal {F}(e_t)\) or \(\mathcal {E}(e_t)\),

if \(S'\) is obtained by a merge, then \(S_1\) is in \(\mathcal {E}',\mathcal {A}_+'\) or \(\mathcal {E}_+'\) and \(S_2\) is in \(\mathcal {C}(e_t)\) or \(\mathcal {E}(e_t)\),

if \(S'\) is obtained with an absorption, then \(S_1\) is in \(\mathcal {A}'\) or \(\mathcal {A}_+'\) and \(S_2\) is in \(\mathcal {C}(e_t)\), and

if \(S'\) is obtained by a closing, then \(S_1\) is in \(\mathcal {A}_+'\) or \(\mathcal {E}_+'\) and \(S_2\) is in \(\mathcal {C}(e_t)\).
Thus, there are 25 complete compositions to consider. If \(S_2 \in \mathcal {C}(e_t)\) (resp. \(\mathcal {E}(e_t)\)) and \(S'\) is obtained by a closing (resp. merge), then \(S'\) is closeable (resp. extensible) if \(S_1\) contains more than one closeable (resp. extensible) alternating path or absorbent, otherwise. Hence, a complete composition obtained with a closing or a merge is not included in a unique set among those defined. This problem can be solved by ignoring certain solutions: \(S'\) can be ignored if there is another solution in \(G_t\) with the same cardinality.

1
Suppose \(S'\) is obtained with a closing (resp. merge) and \(S_1\) contains more than one closeable (resp. extensible) alternating path. Let \(p_1\) and \(p_2\) be closeable (resp. extensible) alternating paths of \(S_1\). There is a solution \(S'_1\) similar to \(S_1\) except that \(p_1\) and \(p_2\) have been closed into a cycle (resp. merged into a unique alternating path) during a previous step. We can obtain a solution in \(G_t\) with the same cardinality as \(S'\) by juxtaposing \(S'_1\) and \(S_2\). Thus, \(S'\) can be ignored, and we suppose that a solution obtained with a closing does not contain a closeable alternating path (i.e. is not in \(\mathcal {A}_+\) or \(\mathcal {E}_+\)). Likewise, we can suppose a solution obtained with a merge between a solution of \(\mathcal {E}'\cup \mathcal {E}_+'\) and a solution of \(\mathcal {E}(e_t)\) does not contain an extensible alternating path (i.e. is not in \(\mathcal {E}(c')\) or \(\mathcal {E}_+\)).

2
Assume that one of the following conditions is true. (1) \(S_1 \in \mathcal {A}_+', S_2 \in \mathcal {E}(e_t)\) and \(S'\) is obtained by a merge, (2) \(S_2 \in \mathcal {E}_+', S_2 \in \mathcal {F}(e_t)\) and \(S'\) is obtained by a merge, (3) \(S_1 \in \mathcal {A}_+', S_2 \in \mathcal {F}(e_t)\) and \(S'\) is obtained by an absorption. Let p be a closeable alternating path of \(S_1\) that is absorbed or merged in \(S'\). There is a solution \(S'_1\) similar to \(S_1\) except that all nonmatching edges of p have been merged or absorbed during previous steps. We can obtain a solution in \(G_t\) with the same cardinality as \(S'\) by juxtaposing \(S'_1\) and \(S_2\). Thus, \(S'\) can be ignored. Fig. 8 shows an example of case (3).
The second item allows us to ignore three complete compositions: there are 22 still to be considered. Each of these complete compositions is in only one of the six sets of solutions among \(\mathcal {F}(c')\), \(\mathcal {A}(c')\), \(\mathcal {E}(c')\), \(\mathcal {A}'_+\) and \(\mathcal {E}'_+\).

Suppose \(S'\) is frozen. The only feasible operation to obtain \(S'\) is juxtaposition because an addition of an edge of \(E(c')\setminus S\) creates an absorbent solution. \(S_1\) and \(S_2\) are frozen as, otherwise, their juxtaposition is not frozen. Thus, line 9 is correct.

Suppose \(S'\) is absorbent. Thus, \(S'\) contains at least one edge in \(\mathcal {E}(c')\setminus S\).

If \(S_2\) is frozen, then the only feasible operation is juxtaposition and \(S_1\) is absorbent.

If \(S_2\) is extensible, then its extensible alternating path is merged with an extensible alternating path of \(S_1\) that is not closeable. Thus, \(S_2\) is in \(\mathcal {E}'\).

If \(S'\) results from an absorption, then \(S_1\) is absorbent and \(S_2\) is closeable.

If \(S'\) results from a closing, then \(S_1\) and \(S_2\) are closeable. Since the resulting solution is absorbent, \(S_1\) is in \(\mathcal {A}_+'\).
Hence, line 10 is correct.


Suppose \(S'\in \mathcal {A}_+\). Then, \(S'\) is extensible and does not contain any extensible alternating paths.

If \(S'\) results from a juxtaposition, then \(S_1\) does not contain an extensible alternating path and \(S_2\) is either frozen or closeable. In the first case, \(S_1\) must be closeable and therefore \(S_1\in \mathcal {A}_+'\). In the second case, \(S_1\) is in \(\mathcal {F}', \mathcal {A}'\) or \(\mathcal {A}_+'\).

If \(S'\) results from a merge, then \(S_1\) is closeable and \(S_2\) is either extensible or closeable. In the first case, the extensible alternating path of \(S_1\) is merged with an extensible alternating path of \(S_2\) so that the resulting solution is not extensible. Thus, \(S_1\) is in \(\mathcal {E}_+'\). In the second case, \(S_1\) does not contain an extensible alternating path since otherwise \(S'\) is extensible. Thus, \(S_1\) is in \(\mathcal {A}_+\).
Hence, line 11 is correct.


Suppose \(S'\) is extensible. Then, either \(S_1\) contains an extensible alternating path or \(S_2\) is extensible.

If \(S_1\) is extensible and \(S'\) results from a juxtaposition, then \(S_2\) is not closeable since otherwise the resulting solution is also closeable. Thus, \(S_2\) is frozen or extensible.

If \(S_1\) is extensible and \(S'\) results from a merge. Then, since we only consider solutions of \(\mathcal {E}'\) with a unique extensible alternating path, \(S_2\) cannot be extensible since otherwise the resulting solution is absorbent. Thus, \(S_2\) is closeable.

If \(S_1\) is in \(\mathcal {E}_+'\), then since \(S'\) is not closeable, the extensible alternating path of \(S_1\) is either merged with an alternating path or closed into a cycle with a closeable alternating path. Thus, \(S_2\) is extensible and \(S'\) results from a merge or \(S_2\) is closeable and \(S'\) results from a closing.

If \(S_2\) is extensible and \(S_1\) does not contain any extensible or closeable alternating path, then \(S'\) results from a juxtaposition and \(S_1\) is frozen or absorbent.
Hence, line 12 is correct.


Suppose \(S'\) is in \(\mathcal {E}_+\). Then, \(S'\) is closeable and contains one extensible alternating path. Recall that we ignore solutions resulting from merge between a solution of \(\mathcal {E}_+\) and a closeable solution. Thus, \(S'\) results from a juxtaposition and either \(S_1\) or \(S_2\) contains an extensible alternating path.

If \(S_1\) is in \(\mathcal {E}_+'\), then \(S_2\) can be any solution.

If \(S_1\) is in \(\mathcal {A}_+'\), then for \(S'\) to contain an extensible alternating path, \(S_2\) must be extensible.

If \(S_1\) is extensible, then for \(S'\) to contain a closeable alternating path, \(S_2\) must be closeable.
Hence, line 13 is correct.

As after these assignments, each of the solutions of \(G_t\) is in a unique set and is a composition of a solution of \(G_{t1}\) and \(G^*(e_t)\), computed values for the table entries are correct for \(G_t\).
Finally, after the execution of the loop, computed values for sets \(\mathcal {F}(c'), \mathcal {A}(c')\) and \(\mathcal {E}(c')\) are correct for \(G_k = G^*(c')\). It remains to compute the value of the table entry for \(\mathcal {C}(c')\). Sets containing closeable alternating paths are exactly the sets \(\mathcal {A}_+\) and \(\mathcal {E}_+\), thus \(\mathcal {A}_+\cup \mathcal {E}_+= \mathcal {C}(c')\). Hence, the assignment line 15 is correct.□
Clique
Let c be a clique of \(G^*\) and let d be the upper door of c. We show in this part how to compute the table entries for the sets \(\mathcal {F}(c)\) and \(\mathcal {E}(c)\). Note that, since the edge between \(G^*(c)\) and its parent is a bridge, the sets \(\mathcal {C}(c)\) and \(\mathcal {A}(c)\) are empty. Let e be the alternating element of c containing the upper door d of c. The idea is to first compute the table entries for the graph \(G^*(e)\) and then merge the obtained table entries to the table entries of the subclique. If e is an alternating path and d is an extremity of e, we replace \(\mathcal {E}(e)\) by two intermediate sets \(\mathcal {\mathcal {E}}_{d}\) and \(\mathcal {\mathcal {E}}_{d'}\). Let \(S'\) be a solution of \(G^*(e)\). Then, \(S' \in \mathcal {\mathcal {E}}_{d}\) if and only if \(S' \in \mathcal {E}(e)\) and d is an extremity of an alternating path of \(S'\). Likewise, \(S' \in \mathcal {\mathcal {E}}_{d'}\) if and only if \(S' \in \mathcal {E}(e)\) and d is not an extremity of an alternating path of \(S'\). Note that \(\mathcal {E}(e)= \mathcal {\mathcal {E}}_{d}\cup \mathcal {\mathcal {E}}_{d'}\). In order to compute these two sets, we reuse the value of \(\mathcal {I}_e\), computed in \(compute\_alternating\_element\).
Lemma 7
For any clique c, the values of the table entries provided by Algorithm 8 are correct for the sets \(\mathcal {F}(c)\) and \(\mathcal {E}(d)\).
Proof
Suppose e is an alternating path and the upper door d of c is an extremity of e. Let \(d'\) be the other extremity of e. First, we compute the table entries for the sets \(\mathcal {C}(e),\mathcal {F}(e),\mathcal {\mathcal {E}}_{d}\) and \(\mathcal {\mathcal {E}}_{d'}\). Suppose that the values of the table entries provided by \(compute\_alternating\_element(p)\) are correct for the sets \(\mathcal {C}(e)\) and \(\mathcal {F}(e)\). It remains to compute the table entries for the sets \(\mathcal {\mathcal {E}}_{d}\) and \(\mathcal {\mathcal {E}}_{d'}\). We recall that \(\mathcal {I}_e\) is the juxtaposition of all frozen solutions of the inner vertices of e.

A solution \(S'\) of \(G^*(e)\) is in \(\mathcal {\mathcal {E}}_{d}\) if and only if \(S'\) is in \(\mathcal {E}(e)\) and no nonmatching edge is incident to d in \(G^*(e)\). Thus, \(\mathcal {\mathcal {E}}_{d}\) is the juxtaposition of e, \(\mathcal {I}_e\), \(\mathcal {F}(d)\) and \(\mathcal {E}(d')\), implying that line 5 is correct.

Similarly, a solution \(S'\) of \(G^*(e)\) is in \(\mathcal {\mathcal {E}}_{d'}\) if and only if \(S'\) is in \(\mathcal {E}(e)\) and no nonmatching edge is incident to \(d'\) in \(G^*(e)\). Thus, \(\mathcal {\mathcal {E}}_{d}\) is the juxtaposition of e, \(\mathcal {I}_e\), \(\mathcal {E}(d)\) and \(\mathcal {F}(d')\), implying that line 6 is correct.
Further, we show that the table entries computed for the set \(\mathcal {F}(c)\) and \(\mathcal {E}(c)\) are correct.

A solution \(S'\) of \(G^*(c)\) is frozen if and only if \(S'\) contains an edge incident to d. This is the case if the subsolution of \(S'\) in \(G^*(e)\) is in \(\mathcal {F}(e)\) or \(\mathcal {\mathcal {E}}_{d'}\) or if S is obtained by a merger operation, an absorption operation or a closing operation. Thus, line 8 is correct.

A solution \(S'\) of \(G^*(c)\) is extensible if and only if S does not contain an edge incident to d. This is the case if the subsolution of S in \(G^*(e)\) is in \(\mathcal {C}(e)\) or \(\mathcal {\mathcal {E}}_{d}\) or if \(S'\) is obtained by a merger operation and \(d'\) is an extremity of an alternating path in the subsolution of S in \(G^*(e)\). Thus, line 10 is correct.
Now, suppose that the upper door d of c is an inner vertex of e. In that case, a subsolution \(S'\) of \(G^*(c)\) is necessarily frozen. Then any feasible composition of a solution of \(G^*(c')\) and a solution of \(G^*(e)\) is a frozen solution and thus, line 13 is correct. Similarly, since no extensible solution of \(G^*(c)\) exist, line 14 is correct.□
Feasibility function
We can now provide an answer to the feasibility of finding a solution for \(\textsc {Scaffolding}\) by using Algorithm 9. Let r be the root of \(G^*\). Notice than since r does not have an upper door then the subclique of r corresponds to r. Thus, it is not possible to call \(compute\_clique\) on r. That is why the first recursive call of the algorithm is made with the function \(compute\_subclique\).
Corollary 2
Given a partial solution S, Algorithm 9 returns true if and only if \((G^*, M^*)\) can be decomposed into \(\sigma _p\) alternating paths and \(\sigma _c\) alternating cycles. The time complexity of the algorithm is \(\mathcal {O}(V(G^*) \cdot \sigma _c^2)\).
Proof
Since \(G^*(root) = G^*\), there is a solution S with \(\sigma _p(S)=\sigma _p\) and \(\sigma _c(S)=\sigma _c\), if and only if S is in \(\mathcal {C}(root)\), \(\mathcal {F}(root)\), \(\mathcal {A}(root)\), or \(\mathcal {E}(root)\). Thus, the return of the function indicates if such a solution exists and then the algorithm is correct. Concerning the time complexity, the composition operations are executable in \(\mathcal {O}(\sigma _c^2)\) time. Thus, without taking into account the recursive calls, the time complexity of Algorithm 5, Algorithm 6, Algorithm 7 and Algorithm 8 in one iteration of a loop is \(\mathcal {O}(\sigma _c^2)\). Let C denote the number of cliques in GG. In Algorithm 5, the number of iterations made by all calls of this function depends on C and then the time complexity of all these iterations is \(\mathcal {O}(C \cdot \sigma _c^2)\). Similarly, we can show that the time complexities of the iterations made by all calls of Algorithm 6, Algorithm 7 and Algorithm 8 are \(\mathcal {O}(V \cdot \sigma _c^2)\), \(\mathcal {O}(M^* \cdot \sigma _c^2)\) and \(\mathcal {O}(C \cdot \sigma _c^2)\). Then, the time complexity of all iterations in all functions is \(\mathcal {O}((V) + M^* + C) \cdot \sigma _c^2)\) and since the number of matching edges and the number of cliques is bounded by the number of vertices of \(G^*\), we have a time complexity \(\mathcal {O}(V(G^*) \cdot \sigma _c^2)\).□
A running example is depicted in Fig. 9 and Example 1 (Tables 1, 2, 3, 4 and 5).
Example 1
Running example on the graph depicted in Fig. 9. Tables 1, 2, 3 and 4 depicte the table entries resulting from Algorithms 5 to 8, respectively. Table 5 display the values of the table entries after each iteration of alternating element for the subclique \(c'_5\). Let c be the value given by the column “#cycles” and x be the item considered in the first column. For each X in \(\mathcal {F} ,\mathcal {A}, \mathcal {A}_+, \mathcal {E}, \mathcal {E}_+\) and \(\mathcal {C}\), the interval given by the column X corresponds to [X(x), c].
Approximation result
We now prove the following approximation result.
Theorem 4
Algorithm 1 provides a solution for \((\sigma _p,\sigma _c)\)scaffolding in connected cluster graphs with an approximation ratio of at most five and a time complexity \(\mathcal {O}(V\cdot E(G^*) \cdot \sigma _c^2)\). The approximation ratio is tight.
Proof We suppose that the input of the algorithm is a scaffold graph \((G^*,M^*,\omega )\) with nonnegative weights and such that \(G^*\) is a path connected cluster graph. We first show that the algorithm is correct. Note that, since each time we add an edge e to S, we remove from E all incident non matching edges to e, the set S induces only paths and cycles.
If it is not possible to build a solution from the graph, then the feasibility condition is not verified and then the algorithm returns an error. Otherwise, since we ensure that the feasibility condition is verified at each step, when the algorithm terminates, then it builds \(\sigma _p\) paths and \(\sigma _c\) cycles.
Now, we prove the approximation ratio. Since they always appear in any solution, we do not consider the edges of \(M^*\) in what follows. Notice that, since there is, for each path, one chosen edge less than the number of involved matching edges, and for a cycle, the same number of chosen edge as the number of involved matching edges, then the number of nonmatching edges in every solution is exactly \(n  \sigma _p\).
We denote by \(e_1,\dots ,e_m\) the edges of the graph \(G^*\), sorted in nonincreasing order by their weights. We denote by \(e^A_1,\dots ,e^A_{n\sigma _p}\) the edges of the solution \(S_A\) given by Algorithm 1, sorted in nonincreasing order by their weights. In the same way, we denote by \(e^{opt}_1,\dots ,e^{opt}_{n\sigma _p}\) the edges of an optimal solution \(S_{opt}\) for the problem, also sorted in nonincreasing order. Both sequences \(e^A_1,\dots ,e^A_{n\sigma _p}\) and \(e^{opt}_1,\dots ,e^{opt}_{n\sigma _p}\) are clearly subsequences of \(e_1,\dots ,e_m\). Let \(\varphi : S_{opt} \rightarrow S_A\) be a mapping such that
Inequality (1) indicates that for each \(e\in E\) in an optimal solution, there is an edge \(\varphi (e) \in S_A\) such that the weight of this latter edge is at least the weight of e. Whereas (2) states that for each \(e \in S_A\), we may associate e to at most four edges of the optimal solution. In the following, we prove that it is possible to define a mapping \(\varphi\) satisfying these inequalities.
The algorithm may decide not to choose an edge \(e^{opt}_i\) for four main reasons:

\(e^{opt}_i\) is eliminated because it is in R, when an edge \(e^A_j\) is chosen. In this case, we have \(\omega (e^A_j) \ge \omega (e_i^{opt})\) because only edges appearing after \(e^A_j\) in the ordered list can be in R. When an edge \(e^A_j\) is chosen, it can eliminate at most two edges of optimal solution by updating of the list of edges (see Fig. 10). We assign \(\varphi (e^{opt}_i) = e^A_j\) in this case. (1) is satisfied by construction, and (2) holds when considering only the optimal edges which are eliminated by this way.

\(e^{opt}_i\) is eliminated because its addition disconnects the graph and the number of alternating cycles and alternating paths required to cover the graph becomes too big. This happens in one of the following two cases.

\(e^{opt}_i\) closes a cycle. In that case, there is at least one edge \(e^A_{j}\) in this cycle, and since it has been chosen before the algorithm considers \(e^{opt}_i\), we necessarily have \(\omega (e^{A}_j) \ge \omega (e^{opt}_i)\). Thus, we assign \(\varphi (e^{opt}_i)= e^A_j\). Then, (1) is satisfied by construction. The edge \(e^A_j\) has been already chosen, may have eliminated at most two optimal edges, but (2) is still satisfied.

\(e^{opt}_i\) closes a door d and one bridge dx incident to d is necessary to construct a solution with the remaining edges. There is a door y which has been closed by an edge \(e^A_j\) in a previous step and this forces dx to be in \(S_A\). Since closing a door increases by at most one the minimum number of alternating paths required to cover the graph, the closing of y forces at most one bridge of \(G^*\) to be in \(S_A\). Thus, the closing of y prevents d and x from closing, that is, at most two edges of \(S^{opt}\), incident to d and x respectively, can be associated to \(e^A_j\) Then, (1) is satisfied by construction. The edge \(e^A_j\) may have eliminated at most two optimal edges in R and may prevent the closing of a cycle, but (2) is still satisfied.


\(e^{opt}_i\) is eliminated because its inclusion would merge two paths \(p_1\) and \(p_2\). If \(e^{opt}_i\) is not a bridge and \(p_1\) and \(p_2\) are a singleedge paths, then the number of alternating cycles and paths are reached in S, that is \(\sigma _c = c\), \(\sigma _p = p\) and \(S = S_A\). Then, we can find an edge \(e^A_j\) such that \(\varphi ^1(e^A_j)=0\) and we assign \(\varphi (e^{opt}_i) = e^A_j\). Then, (1) and (2) are satisfied by construction. Otherwise, the algorithm eliminates \(e^{opt}_i\) because one of the merged paths must be closed into a cycle to reach the correct number of alternating cycles. Otherwise, there is an edge \(e^A_j\) in \(S_A\) considered before \(e^{opt}_i\) in the algorithm such that \(\varphi ^1(e^A_j) \le 3\) (since otherwise the path would be already closed into a cycle) and then we assign \(\varphi (e^{opt}_i) = e^A_j\). Again, (1) and (2) are satisfied by construction.
From the previous discussion and by (1) and (2), clearly we have:
The ratio is tight, as shown by the example depicted in Fig. 11.
Concerning the complexity, the edges can be sorted in \(\mathcal {O}(V(G^*) \log E(G^*))\) time. The feasibility function is called \(E(G^*)\) times. Thus, the time complexity of the algorithm is \(\mathcal {O}(E(G^*) \cdot V(G^*) \cdot \sigma _c^2)\).
Experimental results
In this section, we compare the performance of Algorithm 1 with three different feasibility functions and an integer linear programming formulation [15] implemented with ILOG CPLEX [16].
Dataset
We reuse the dataset already used in [9], which was obtained with the following pipeline:

1
Choice of a reference genome, for instance on the nucleotide database from NCBI^{Footnote 2}. Table 6 presents selected genomes used for our experiments. We chose a panel of genomes of various origins and sizes.

2
Simulation of pairedend reads, using wgsim [17]. The chosen parameters are an insert size of 500bp and a read length L of 100bp.

3
Assembly using the de novo assembly tool, based on a De Bruijn graph efficient representation: minia [18] with kmer size \(k=30\).

4
Mapping of reads on contigs, using bwa [19]. This mapping tool was chosen according to results obtained by Hunt [20], a survey on scaffolding tools.

5
Generation of scaffold graph from the mapping file.
Statistics on the numbers of vertices and edges in produced scaffold graphs can be viewed in Table 7.
Feasibility functions
There is no polynomialtime computable feasibility function in the general case. Thus, to use the greedy algorithm with a specific feasibility function on a real instance, we must transform it. For this, we construct a supergraph by adding edges of weight zero. We compare three feasibility functions, defined on complete graphs, connected cluster graphs and block graphs^{Footnote 3}, respectively. Note that the construction of a complete supergraph requires the largest amount of edge additions whereas the least amount of edge additions is required for the construction of a block supergraph. We already showed in [9] that the computed ratio is close to one on real instances, that is, relatively far from the theoretical ratio of 3. The aim of these experiments is to answer the two following questions:

Can greedy algorithms on connected cluster graphs and block graphs be used on large scaffold graphs, and what is its associated computation time?

Do we get a better practical ratio if the amount of additional edges is smaller (e.g. the completion rate, see Table 7, is smaller)? In other words, do we obtain better results on block graphs and connected cluster graphs than in complete graphs?
Results
Experiments were run on a personal computer with four i7 processors at 1.9GHz and 16GB RAM. Memory usage was very light, even on the biggest instance anopheles. Table 8 shows scores and computation times for every instance. We can see that greedy computation times are less than few seconds except for anopheles, where the connected cluster graph version and the block graph version need a few minutes. As expected, the greedy algorithms are much faster than the ILP formulation in every case. These results let us answer to our first question: connected cluster graph and block graph versions of the greedy algorithm are capable of treating big instances, however the computation time is significantly bigger than the complete version. Concerning the scores, we can see that the three greedy algorithms have the same score for most of the data. The connected cluster graph and block graph versions have a slightly better score in four instances: anopheles, anthrax, sacchr3 and sacchr12. Moreover, connected cluster graph and block graph versions have the same score in all instances except in anopheles, where the block graph version improves the score of the connected cluster graph version by three (which is not really significant compared to the absolute values). These results indicate that the answer to the second question is positive. However, the differences between scores are not significant enough to be completely affirmative. We can think that using the greedy algorithm with feasibility function defined on a sparser class of graphs may lead to better results.
Conclusion and future work
We presented in this paper the first polynomialtime algorithm approximating the scaffolding problem on noncomplete graphs. Using a dynamic programming approach, we exploited the treelike nature of connected cluster graphs to extend the feasibility function and the analysis of the approximation ratio. We also showed that this new algorithm provides slightly better results on real data than the greedy algorithm on complete graphs, although its theoretical ratio is worse. This leads us to the hypothesis that using a feasibility function defined on a graph class close to the original instance produces better results. This is surprising since, intuitively, algorithms on superclasses can choose from a larger set of edges to build solutions (any solution on the more restricted class is also a solution in the more general class). A natural extension of this work is to consider sparser graphs: for example, one could replace cliques in connected cluster graphs by cobipartite graphs as the feasibility function is polynomialtime computable in this case [8]. One may also explore the possibility of exploiting randomized algorithms to improve the ratio [6].
Availability of data and materials
The datasets used and/or analysed during the current study are available from the corresponding author on reasonable request.
Notes
Note that v follows u in the genomic sequence if and only if u follows v in its reverse complement. Therefore, scaffolds are modeled as undirected graphs in this work.
A block graph is a graph in which every biconnected component is a clique (note that a connected cluster graph is a special case of block graph).
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TD and RF developed complexity and inapproximability results. TD, AC, RG and MW conceived the feasibility function on connected cluster graphs. TD implemented the greedy algorithm and performed the tests. All authors read and approved the final manuscript.
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Davot, T., Chateau, A., Fossé, R. et al. On a greedy approach for genome scaffolding. Algorithms Mol Biol 17, 16 (2022). https://doi.org/10.1186/s1301502200223x
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DOI: https://doi.org/10.1186/s1301502200223x
Keywords
 Genome scaffolding
 Complexity
 Approximation
 Dynamic programming
 PolyAPXhardness