When a given square sparse matrix can be permuted into a lower triangular form?

For some, the question in the title of this post could look trivial. After all, they have been doing LU decomposition in Matlab and using Matlab’s mldivide (or “backslash”) to solve systems with the factor matrices. These are usually done without worrying if the pivoting took place and permuted those factor matrices (Matlab used to call them psychologically triangular matrices. For “why?” and “who said that?”, see here). Because, the help page of mldivide says that it tests if a matrix “is square, is not diagonal, looks triangular, is actually triangular, there are zeros in the diagonal, is a permuted triangular”. So this must be taken care of…

Davis’s “Direct Methods for Sparse Linear Systems” book (a real gem) asks in its Exercise 3.7:

…Consider a matrix \mathbf{A} that may be a permuted upper or lower triangular matrix with both rows and columns permuted by unknown permutations \mathbf{P} and \mathbf{Q}. Write an algorithm that determines if the matrix is in this form and, if so, solves \mathbf{A}x = b….

Then the book hints an algorithm, which builds the permutation matrices essentially by picking a row with a single nonzero, ordering that row and the associated column, and then by deleting both from the matrix and so on. But this is ok, only if we can recover a full diagonal! What if, at some point, we have rows or columns with no entries at all? The “algorithm” does not handle this case.

It is perhaps ok in the context of LU decomposition, in which case one is warned of about the numerical rank deficiency during LU decomposition (before one attempts to solve the linear systems). But in general it is not ok; at Mathworks, folks are aware of this (see here). If the matrix is not coming from LU decomposition, then we may well be in trouble. We are back to the question: When a given square sparse matrix can be permuted into a lower triangular form?

Some other cases are also easy. If there is a perfect matching in the given matrix, then all standard matching algorithms find it quickly, and after permuting the matching entries to the diagonal, we can check if the matrix is reducible into 1\times 1 blocks (or the directed graph of \mathbf{B}=\mathbf{AQ} is acyclic, where \mathbf{Q} permutes the columns of \mathbf{A} to have a zero-free diagonal in \mathbf{B}). Every step of this is linear time—even computing the matching with the standard tricks). So this case is handled. The question is complicated when we do not have a perfect matching. It is in fact NP-complete to decide if a given square sparse matrix can be permuted into a lower/upper triangular form [1] with unsymmetric permutations.

So writing it once more to emphasize it:

It is easily decidable to test if \mathbf{A} can be symmetrically permuted to a lower/upper triangular form. The directed graph of \mathbf{A} should be acyclic (or \mathbf{PAP}^T is lower/upper triangular).

It is NP-complete to decide if a given square sparse matrix can be permuted to a lower/upper triangular form. Otherwise said, one can find two permutation matrices \mathbf{P} and \mathbf{Q} such that \mathbf{PAQ} is lower/upper triangular only with non-polynomial time algorithms, assuming P\neq NP.

The topic was discussed elsewhere before.


  1. Guillaume Fertin, Irena Rusu, and Stéphane Vialette:
    Obtaining a triangular matrix by independent row-column permutations. ISAAC 2015: 165–175 (doi)

Karp–Sipser heuristic:

On the degree-1 and degree-2 reduction rules

Karp–Sipser (KS) [5] heuristic is a well-known method to obtain high quality matchings in undirected graphs, bipartite graphs, and hypergraphs.  This heuristic is based on two reduction rules. The first rule says that if there is a vertex with degree one, then we can match that vertex with its unique neighbor and remove both from the graph without any loss.  The second rule says that if there is a vertex v with degree two, then its neighbors can be coalesced, and the vertex v can be discarded. A maximum matching on the reduced graph can be extended to a maximum matching on the original graph. If there is no degree-1 or degree-2 vertices, then a randomly selected edge is used to match two vertices, upon which both of them are removed.

The (full) heuristic with the degree-2 reduction rule turns out to be difficult to analyse. That is why most studies are conserved with the variant with the degree-1 reduction rule. This variant has excellent practical performance [6,7], good theoretical guarantees for certain classes of random graphs [2,5], and can be made to have an approximation ratio around 0.866 [3] for bipartite graphs and a close-by ratio on general graphs [4]. There are graphs for which KS with only the degree-1 reduction can obtain worse results. The following is an example and a table from [3].

The bipartite graph instances giving hard time to KS with the degree-1 rule correspond to the following matrices (rows form one part, and the columns form the other part) . The rows and columns are split into two sets in the middle, yielding a 2\times 2 structure. The (1,1)-block is full, the (2,2)-block is empty, and there is an identity matrix at the other two blocks. In order to hide the perfect matching composed of those two identity matrices, additional nonzeros are added to the last t rows in the first row block, and the last t columns in the first column block so that those rows and columns become dense. For n=3200, and with different thickness t\in \{2,4,8,16, 32\} in the full rows, KS with only the degree-1 rule obtains the performance shown in the table.


t Performance
2 0.782
4 0.704
8 0.707
16 0.685
32 0.670

We want to now turn attention to the full KS (that is with the degree-1 and degree-2 reduction rules). Anastos and Frieze [1] state in the abstract of a recent arXiv paper:

In a seminal paper on finding large matchings in sparse random graphs, Karp and Sipser (FOCS 1981) proposed two algorithms for this task. The second algorithm has been intensely studied, but due to technical difficulties, the first algorithm has received less attention. Empirical results in the orignal paper suggest that the first algorithm is superior. In this paper we show that this is indeed the case, at least for random cubic graphs. We show that w.h.p. the first algorithm will find a matching of size n/2-O(\log n) on a random cubic graph (indeed on a random graph with degrees in \{3, 4\}). We also show that the algorithm can be adapted to find a perfect matching w.h.p. in O(n) time, as opposed to O(n^{3/2}) time for the worst-case.

This wonderful result shows improvement with the degree-2 reductions (over using only degree-1 reductions): instead of leaving out O(n^{1/5}) vertices as KS with the degree-1 rule only, it leaves out only \log n vertices. We can also create instances for which KS with both reduction rules obtains perfect matchings, whereas that with the degree-1 reduction obtains inferior results. These instances are upper (or lower) Hessenberg matrices, or more sparser versions of it (for example if we discard all but the first and the last subdiagonal nonzeros). A figure is shown below with only two subdiagonal nonzeros. On these instances, an immediate application of the degree-2 reduction triggers a chain of degree-1 reductions, until we have a 2\times 2 matrix, whose perfect matching is obtained by a degree-2 and then a degree-1 reduction. On the other hand, KS with only the degree-1 rule obtains the performance shown in the table for different n (the performance reported in the table is average of 10 different random runs on the same instance).


n Performance
500 0.828
1000 0.790
5000 0.754

There is a catch though. While there is a linear time, straightforward implementation of the degree-1 rule, the straightforward implementation of the degree-2 rule can be quadratic in the worst case; consider repetitive merges with the same vertex which will happen with an Hessenberg matrix. We can perhaps assume that the worst cases are rare in practice, but nonetheless a worst-case quadratic run time complexity is a lot!


  1. Michael Anastos and Alan Frieze, Finding perfect matchings in random cubic graphs in linear time, arXiv preprint arXiv:1808.00825, Nov., 2018. (url)
  2. Jonathan Aronson, Alan Frieze, and Boris G. Pittel, Maximum matchings in sparse random graphs: Karp-Sipser revisited, Random Structures & Algorithms, 12 (1998), pp. 111-177.
  3. Fanny Dufossé, Kamer Kaya, and Bora Uçar, Two approximation algorithms for bipartite matching on multicore architectures, Journal of Parallel and Distributed Computing, 85 (2015), pp. 62–78. (doi)
  4. Fanny Dufossé, Kamer Kaya, Ioannis Panagiotas, and Bora Uçar, Approximation algorithms for maximum matchings in undirected graphs, Proceedings of the Seventh SIAM Workshop on Combinatorial Scientific Computing, Bergen, Norway, 2018, pp. 56–65. (doi)
  5. Richard M. Karp and Michael Sipser, Maximum matching in sparse random graphs, 22nd Annual IEEE Symposium on Foundations of Computer Science (FOCS), 1981, pp. 364–375. (doi)
  6. Johannes Langguth, Fredrik Manne, and Peter Sanders, Heuristic initialization for bipartite matching problems, Journal of Experimental Algorithmics, 15 (2010), pp. 1.1–1.22. (doi)
  7. Jakob Magun, Greedy matching algorithms, an experimental study, Journal of Experimental Algorithmics, 3 (1998). (doi)

Recent news on the minimum fill-in problem

We have just seen a a paper in the Proceedings of the Twenty-Eighth Annual ACM-SIAM Symposium on Discrete Algorithms (SODA), 2017. The proceedings is available online for the curious. A paper from the proceedings is the subject of this post.

Minimum fill-in: Inapproximability and almost tight lower bounds,
by, Yixin Cao and R. B. Sandeep (url).

The minimum fill-in problem is one of the core problems in sparse direct methods. This problem is NP-complete [6]. The NP-completeness of the problem was first conjectured to be true in 1976 by Rose, Tarjan, and Lueker [5] in terms of the elimination process on undirected graphs. Then Rose and Tarjan [4] proved in 1978 that finding an elimination ordering on a directed graph that gives minimum fill-in is NP-complete (there was apparently a glitch in the proof which was rectified by Gilbert [1] two years later). Finally, Yannakakis [6] proved the NP-completeness of the minimum fill-in problem on undirected graphs in 1981.

The sparse matrix community needs a heuristic to tackle the minimum fill-in problem, as direct methods are used in many applications. There are quite efficient and effective heuristics for this purpose, which are variants of the approximate minimum degree and incomplete nested dissection algorithms. These heuristics are efficiently implemented in programs that are wide spread. The mentioned heuristics do not have any performance guarantees (except for classes of graphs such as graphs with good separators) but they are very-well established. So there is an NP-complete problem, practitioners have some heuristics and are happy with what they have.

On the other hand, the minimum fill-in problem poses many challenges to the theoreticians. Natanzon, Shamir, and Sharan [2] obtained an algorithm that approximates the minimum fill-in (k) by bounding the fill by 8k^2 in O(k n^3) time, for a graph with n vertices. The paper by Cao and Sandeep shows that unless P=NP, there is no polynomial time approximation scheme (PTAS) for the minimum fill-in problem. A PTAS is an algorithm which takes an instance of an optimization problem and a parameter \epsilon > 0 and, in polynomial time, produces a solution that is within a factor 1 + \epsilon of the optimum. This is a bad news in a way: no simple heuristic with a performance guarantee for theoretically oriented practitioners is in view.

The paper by Cao and Sandeep contains two other theorems, ruling out algorithms with approximation 1+\epsilon and with a running time 2^{O(n^{1-\delta})}, and exact algorithms with running time in 2^{O(k^{1/2-\delta})}\cdot n^{O(1)}, assuming exponential time hypothesis (ETH). Here n is again the number of vertices, k is an integer parameterizing the fill-in, and \delta is a small positive constant. ETH posits that the satisfiability problem with at most three variables per clause cannot be solved in 2^{o(p+q)} time, where p and q denote the number of clauses and variables, respectively, in the problem.

If you are interested in these problems, then see the tree-width and minimum fill-in challenges. The minimum tree-width is related to the shortest elimination tree over all elimination orderings of an undirected graph, which one of us had proved to be NP-complete [3].


  1. John R. Gilbert, A note on the NP-completeness of vertex elimination on directed graphs, SIAM Journal on Algebraic and Discrete Methods, 1 (1980), pp. 292–294 (link).
  2. Assaf Natanzon, Ron Shamir, and Roded Sharan, A polynomial approximation for the minimum fill-in problem, SIAM Journal on Computing, 30, (2000), pp. 1067–1079 (link).
  3. Alex Pothen, The complexity of optimal elimination trees,
    Tech. Report, Department of Computer Science, Penn State University, 1988 (link).
  4. Donald J. Rose and Robert E. Tarjan, Algorithmic aspects of vertex elimination in directed graphs, SIAM Journal on Applied Mathematics, 34 (1978), pp. 176–197 (link).
  5. Donald J. Rose, Robert E. Tarjan, and George S. Lueker, Algorithmic aspects of vertex elimination on graphs, SIAM Journal on Computing, 5 (1976), pp. 266–283 (link).
  6. Mihalis Yannakakis, Computing the minimum fill-in is NP-complete, SIAM Journal on Algebraic and Discrete Methods, 2 (1981), pp. 77–79 (link).

After CSC16 (cont’)

Umit V. Çatalyürek

Umit V. Çatalyürek presented our work on the directed acyclic graph (DAG) partitioning problem. In this problem, we are given a DAG G=(V,E) and an integer k\geq 2. The aim is to partition the vertex set V into k parts V_1,\ldots, V_k in such a way that the parts have (almost) equal weight and the sum of the costs of all those arcs having their endpoint in different parts minimized. Vertices can have weights, and the edges can have costs. Up to now, all is standard. What is not standard is that the quotient graph of the parts should be acyclic. In other words, the directed graph G'=(V', E'), where V'=\{V_1,\ldots,V_k\} and V_i\rightarrow V_j\in E' iff v_i\rightarrow v_j\in E for some v_i\in V_i and v_j\in V_j, should be acyclic.

John R. Gilbert

John R. Gilbert wanted to understand the complexity of the problem, with respect to the undirected version. He is an expert on the subject matter (see, e.g., [2]). He asked what happens if we orient the edges of the n\times n model problem. If you are not familiar with this jargon, it is the n\times n mesh with each node being connected to its immediate neighbor in the four main directions, if those neighbors exist. See the small image for an 8\times 8 example.

8\times 8 model problem.

Partitioning these kind of meshes is a very hard problem. Gary Miller had mentioned their optimal partitioning in his invited talk (about something else). Rob Bisseling [Ch. 4, 1] has a great text about partitioning these meshes and their cousins in 3D. I had surveyed known methods in a paper with Anaël Grandjean [3]. In particular, Anaël found about discrete isoperimetric problems, showed that the shape of an optimal partition at a corner, or inside the mesh was discussed in the related literature. He also showed that the Cartesian partitioning is optimal for edge cut.  Anaël also proposed efficient heuristics which produced connected components. See the cited papers for nice images. Our were based on our earlier work with Umit [4].

Anyhow,  let’s return back to acyclic partitioning of DAGs, and John’s question. He suggested that we should look at the electrical spiral heater to get an orientation. This orientation results in a total order of the vertices. The figures below show the ordering of the 8\times 8 and the 16\times 16 meshes. Only some of the edges are shown; all edges of the mesh, including those that are not shown are from the lower numbered vertex to the higher numbered one.

As seen in the figures above, the spiral ordering is a total order and there is only one way to cut the meshes into two parts with the same number of vertices; blue and red show the two parts.

Theorem: Consider the n\times n mesh whose edges are oriented following the electrical spiral heater ordering. The unique acyclic cut with n^2/2 vertices in each side has n^2-4n+3 edges in cut, for n\geq 8.

The theorem can be proved by observing that the blue vertices in the border (excluding the corners) has one arc going to a red vertex; those in the interior, except the one labeled n^2/2 has 2 such arcs;  the vertex labeled n^2/2 has three such arcs. The condition n\geq 8 comes from the fact that we assumed that there are blue vertices in the interior of the mesh. This is a lot of edges to cut!


Later, John said that he visited the Maxwell Museum of Anthropology at UNM after the CSC16 workshop, and saw that similar designs by the original native New Mexicans.


  1. Rob H. Bisseling, Parallel Scientific Computation: A Structured Approach using BSP and MPI, 1st ed, Oxford University Press, 2004.
  2. John R. Gilbert, Some nested dissection order is nearly optimal. Inf. Process. Lett. 26(6): 325–328 (1988).
  3. Anaël Grandjean and Bora Uçar, On partitioning two dimensional finite difference meshes for distributed memory parallel computers. PDP 2014: 9–16.
  4. Bora Uçar and Umit V. Çatalyürek, On the scalability of hypergraph models for sparse matrix partitioning. PDP 2014: 593–600.
  5. Da-Lun Wang and Ping Wang, Discrete isoperimetric problems, SIAM Journal on Applied Mathematics, 32(4):860–870 (1977).

On HPC Days in Lyon

Last week (6–8 April, 2016) we had an incredible meeting called HPC Days in Lyon. This three day event featured only invited long talks and invited mini-symposia talks. The meeting was organized with generous support of Labex MILYON.

Rob Bisseling and Alex Pothen contributed to a mini-symposium on combinatorial scientific computing.

Rob Bisseling

Rob talked about hypergraph partitioning and how to use it with an iterative solver. We usually get this question: how many mat-vecs (or iterations in a solver) one needs to perform to offset the use of hypergraph partitioning. Rob’s main point in this talk was that one can estimate the number of iterations and spend some more time partitioning the hypergraph, if the number of iterations allow it. He has an ongoing project of optimally bisecting sparse matrices (see the link); his talk included updates (theoretical and practical) to this project. He says he adds a matrix a day to this page. As of now, there are 263 matrices. Chapeau! as the French say.

Also, he said (well maybe slipped out after a few glasses of Côtes du Rhône) that the new edition of his book (Parallel Scientific Computation: A Structured Approach using BSP and MPI) will be coming out.  There are new materials; in particular a few sections on sorting algorithms and a complete chapter on graph algorithms (mainly matching). Stay tuned! Rob will be at SIAM PP next week. I will try to get more  information about his book.

[I have just realized that I did not put Alex’s photo anywhere yet. So let’s have his face too.]

Alex Pothen

Alex discussed approximation algorithms to matching and b-matching problems. He took up the challenge of designing parallel algorithms for matching problems, where the concurrency is usually limited. He discussed approximation algorithms with provable guarantees and great parallel performance for the b-matching and a related edge cover problem. He also discussed an application of these algorithms in a data privacy problem he has been working on.

Alex arrived earlier to Lyon and we did some work. With Alex, we always end up discussing matching problems. This was not exception. We looked at the foundations of bottleneck matching algorithms. Alex and I will be attending SIAM PP16 next week. If you know/like these algorithms, please attend CSC mini-symposia so that we can talk.

I had chaired an invited talk by Yousef Saad.

Yousef Saad

The talk was for 90 minutes, without break! His talk was very engaging and illuminating. I enjoyed very much and appreciated how he communicates deep math to innocent (or ignorant;)) computer scientists. His two books Iterative methods for sparse linear systems (2nd edition) and Numerical Methods for Large Eigenvalue Problems (2nd Edition) are available at his web page and attest this.
Here is a crash course on Krylov subspace methods from his talk.

Let x_0 be an initial guess and r_0=b-Ax_0 be the initial residual.
Define K_m=\textrm{span}\{r_0, Ar_0,\ldots,A^{m-1}r_0\} and L_m another subspace of dimension m.
Basic Krylov step is then:
x_m=x_0 + \delta where \delta\in K_m and b-Ax_m \perp L_m.

At this point, the reader/listener gets the principle and starts wondering what are the choices of L_m that make sense? How do I keep all m vectors? How do I get something orthogonal to them? Yousef had another slide:

1. L_m=K_m; class of Galerkin or orthogonal projection methods (e.g., CG), where \|x^*-\tilde{x}\|=\min_{z\in K_m}\|x^*-z\|_{A}.
2. L_m=AK_m; class of minimal residual methods (e.g., ORTHOMIN, GMRES) where \|b-A\tilde{x}\|_2=\min_{z\in K_m}\|b-Az\|_2.

So we learned the alternatives for L_m, and we probably guessed correctly that we do not need to keep all m vectors in all cases (e.g., CG), sometimes need all (e.g., GMRES without restart), and even if we need we can short-cut and restart. Getting orthogonal vectors could be tougher, especially if we do not store all the m vectors. Now that we have a guide, a feeling, and a few questions we can turn to resources to study.

On the Birkhoff-von Neumann decomposition

Michele Benzi, Alex Pothen and I have been making use of the celebrated Birkhoff-von Neumann theorem on doubly stochastic matrices. The theorem says that any doubly stochastic matrix can be written as a convex combination of a finitely many permutation matrices. Formally, let \mathbf{A} be a doubly stochastic matrix. Then,

\mathbf{A}=\sum_{j=1}^k \alpha_j \mathbf{P}_j\;,

where \alpha_j>0, \sum_j \alpha_j=1 and each \mathbf{P}_j is a permutation matrix.

Given this formulation, one wonders if the decomposition is unique. Well, the answer is “No”. Then, one asks what can be said about the number k. And this is the main topic of this post.

Richard A. Brualdi [1] discusses many things among which a lower bound and an upper bound on k. The minimum number of permutation matrices is equal to the maximum cardinality of a set of nonzeros positions of \mathbf{A} no two of which could appear together in a single permutation matrix in the pattern of \mathbf{A}. An easy lower bound is then equal to the maximum number of nonzeros in a row or a column of \mathbf{A}. The upper bound is \mathrm{nnz}(\mathbf{A})-2n+2, for a fully indecomposable matrix \mathbf{A} with \mathrm{nnz}(\mathbf{A}) nonzeros; more generally if there are b fully indecomposable blocks, then the upper bound is \mathrm{nnz}(\mathbf{A})-2n+b+1.

What about the minimum number of permutation matrices? It turns out that this is an NP-complete problem. Let’s state it in the form of a standard NP-completeness result.

Input:  A doubly stochastic matrix \mathbf{A}.
Output: A Birkhoff-von Neumann decomposition of \mathbf{A} as \mathbf{A} = \alpha_1\mathbf{P}_1 + \alpha_2\mathbf{P}_2 + \cdots +\alpha_k\mathbf{P}_k.
Measure: The number k of permutation matrices in the decomposition.

The NP-completeness of the decision version of this problem is shown in a recent technical report [2].


  1. Richard A. Brualdi, Notes on the Birkhoff algorithm for doubly stochastic matrices, Canadian Mathematical Bulletin, 25(2), 191–199, 1982 (doi).
  2. Fanny Dufossé and Bora Uçar, Notes on Birkhoff-von Neumann decomposition of doubly stochastic matrices, Technical Report, RR-8852, Inria Grenoble Rhône-Alpes, 2016 (link).