Someone asked me today (or yesterday, depending on whose time zone you go by) how to force a group of variables in an optimization model to take the same sign (all nonpositive or all nonnegative). Assuming that all the variables are bounded, you just need one new binary variable and a few constraints.

Assume that the variables in question are $x_1,\dots,x_n$ and that they are all bounded, say $L_i \le x_i \le U_i$ for $i=1,\dots,n$. If we are going to allow variables to be either positive or negative, then clearly we need $L_i < 0 < U_i$. We introduce a new binary variable $y$ and, for each $i$, the constraints$$L_i (1-y) \le x_i \le U_i y.$$If $y$ takes the value 0, every original variable must be between its lower bound and 0 (so nonpositive). If $y$ takes the value 1, every original variable must be between 0 and its upper bound (so nonnegative).

Note that trying to enforce strictly positive or strictly negative rather than nonnegative or nonpositive is problematic, since optimization models abhor strict inequalities. The only work around I know is to change "strictly positive" to "greater than or equal to $\epsilon$" for some strictly positive $\epsilon$, which creates holes in the domains of the variables (making values between 0 and $\epsilon$ infeasible).

## Friday, September 21, 2018

## Wednesday, September 12, 2018

### Choosing "Big M" Values

I seem to bring up "big M" models a lot, so apologies if I end up repeating myself in places here. Not long ago, someone passed along highlights of a "big M" type model to me and asked if he could somehow reformulate to get rid of $M$. I did not see any good way to do it, but I did offer up suggestions about choosing values for $M$, and I thought that might make a decent blog post.

Just for clarity, what I'm referring to is an integer or mixed-integer program (I'll stick to linear objective and constraints) in which binary variables are used to, in loose terms, turn constraints on or off. So a representative constraint might look like the following:$$\sum_i a_i x_i \le b + M(1-y)$$with the $a_i$ coefficients, the $x_i$ variables (discrete, continuous or a mix), $y$ a binary variable and $M$ a "large" coefficient. Think of $M$ as a stunt double for infinity. The notion is that if $y=1$, the constraint is "turned on" and reduces to$$\sum_i a_i x_i \le b.$$If $y=0$, the constraint is "turned off" and reduces to$$\sum_i a_i x_i \le b + M \approx \infty,$$which should be satisfied by any choices for $x$ the solver might make. There are other variations of "big M" constraints, but I'll stick with this one for illustrative purposes.

Choosing a value for $M$ can be a tricky business. Suppose first that we choose $M$ small enough that, when $y=0$ and the constraint should be "off",$$\sum_i a_i x_i^* \gt b + M$$for some solution $x^*$ that

Since it is frequently not obvious how large $M$ needs to be in order to be guaranteed correct, people tend to err on the side of caution by choosing really massive values for $M$. That brings with it a different set of problems (ones often ignored in introductory text books). First, branch-and-bound (or branch-and-cut) solvers tend to rely on the continuous relaxation of subproblems (dropping integrality constraints but keeping the functional constraints). Large values of $M$ make for weak relaxations (producing very loose bounds).

Suppose, for instance, that we are solving a model that both builds roads and routes traffic along those roads. $y$ represents the decision to build a particular road ($y=1$) or not ($y=0$). If we build the road, a cost $c$ is incurred (represented by the term $cy$ in the objective function) and we can send unlimited traffic along it; if not, there is no cost but no traffic is committed. In our constraint, the left side is traffic on the road and $b=0$, so traffic there can either be up to $M$ (if built) or 0 (if not built). Now suppose, for example, that the user chooses $M=10^9$ and the solver, in a continuous relaxation of some subproblem, sets $y=10^{-6}$. The solution to the relaxed problem pays only one millionth of $c$ for the privilege of allowing 1,000 units of traffic on this route, which basically allows it to "cheat" (and perhaps grossly underestimates the cost of any actual solution).

A related problem is limited arithmetic precision. Since double-precision floating-point arithmetic (the way the computer does arithmetic with real numbers) has limited precision, and is thus susceptible to rounding errors, solvers have to establish standards for "close enough to feasible" and "close enough to integer". Continuing the previous example, it is entirely possible that the solver might look at $y=10^{-6}$, decide that is within integrality tolerance, and treat it as $y=0$, possibly leading to what it thinks is an "optimal" solution with 1,000 units of traffic running over a road that does not exist. Oops.

Finally, note that $M$ is part of the coefficient matrix. Mixing very large coefficients (like $M$) with much smaller coefficients can create numerical instability, leading the solver to spend more time computing linear program pivots and possibly leading to totally erroneous solutions. That's too big a topic to get into here, but I'm pretty sure I've mentioned it elsewhere.

Despite all this, "big M" models are still very common in practice. There is a nice paper by Codato and Fischetti [1] that shows how a form of Bender's decomposition can be used to get rid of the $M$ coefficients. I've used it successfully (for instance, in [2]), but Bender's decomposition is an advanced technique (i.e., not for the faint of heart), and is not always supported by high-level modeling languages.

So, if we are stuck with "big M" models, what can we do to choose values of $M$ that are both correct and and

Most "big M" models have more than one constraint (and frequently many) containing a large coefficient $M$. One of my pet peeves is that authors of text books and journal articles will frequently just us $M$, with no subscripts, everywhere such a coefficient is needed. This, in turn, breeds a tendency for modelers to choose one large value for $M$ and use it everywhere.

Back when manuscripts were produced on typewriters, it was a bit of a pain in the ass to put in subscripts, so I can see how the trend would have started. (Quick tangent: Nate Brixius recently did a blog post with a picture of a typewriter, for those too young to be familiar with them. I'll link it here for the benefit of younger readers ... and also note that the one pictured is much more modern than the one I learned on, which was not electric.) Today, when people routinely use LaTeX to write manuscripts, there's not much excuse for omitting one measly subscript here and there. Anyway, my point is that it is usually better to use a different, customized value of $M$ for each constraint that needs one.

In some cases, model context will provide you an obvious choice for $M$. For example, suppose you are selecting warehouse locations and planning shipments from warehouses to customers. A typical "big M" constraint will look like the following (slightly different from our previous constraint but clearly related:$$\sum_j x_{ij} \le M_i y_i.$$Here variable $x_{ij}$ indicates the amount shipped from warehouse $i$ to customer $j$, binary variable $y_i$ is 1 if we use that warehouse and 0 if we do not, and the intent of the constraint is to say that if we do not use (and pay for) a warehouse, we cannot ship anything from it. One obvious choice for $M_i$ is the capacity of the warehouse. A better choice might be the smaller of that capacity or the maximum volume of customers demands that might plausibly be assigned to that warehouse. The latter might be based, for example, on knowing that the company would not ship anything to customers outside a certain distance from the warehouse.

In other cases, there may be some less obvious way to extract suitable (correct and hopefully tight) values for the $M_i$. A while back, I was working on mixed-integer programming models for two group discriminant analysis, and in one paper ([3]) I needed to pick values for $M$. Without going into gory details, I was able to normalize the coefficients of the (linear) discriminant function under construction and then compute valid coefficients $M_i$ by looking at the euclidean distances between pairs of observations. I don't claim that the $M_i$ I came up with were the tightest possible, but they were correct and they produced faster solution of the models than what I got with some arbitrarily large values I initially tried.

Finally, you can actually solve subproblems to get correct and (hopefully) tight $M$ values. Whether this is feasible depends on how many you need and how large and scary your model is. Going back to my first example, the trick would work something like this. Relax all integrality constraints. Drop the constraint in question. If you have any other "big M" constraints for which you have not yet computed tight values of $M$, pick something safely large for those coefficients, trying to avoid numerical problems but not worrying about tightness. Now switch the objective to maximizing $\sum_i a_i x_i - b$. The optimal objective value is your value for $M$. It is clearly correct: any feasible solution to the MIP model is a feasible solution to the LP relaxation, and so the value of $\sum_i a_i x_i - b$ cannot exceed your choice of $M$ (regardless of whether $y$ is 0 or 1). Repeat for each additional constraint containing a "big M" coefficient (switching from minimizing to maximizing if the nature of the constraint warrants it).

The process may be time-consuming, but at least you are solving LPs rather than MIPs. It is also a bit trickier than I made it sound, at least when multiple $M_i$ are involved. You have to guess values for any $M_i$ for which you have not yet solved, and guessing big values for them may result in a looser relaxation than necessary, which in turn may result in an inflated value for the $M_i$ you are currently choosing. You should definitely get correct choices for the $M$ coefficients; it's just the tightness that is in question. Even so, the values you get for the $M_i$ just might be tight enough to save you more time in the MIP solution than it costs to solve the LPs.

[1] Codato, G. and Fischetti, M.

[2] Bai, L. and Rubin, P. A.

[3] Rubin, P. A.

Just for clarity, what I'm referring to is an integer or mixed-integer program (I'll stick to linear objective and constraints) in which binary variables are used to, in loose terms, turn constraints on or off. So a representative constraint might look like the following:$$\sum_i a_i x_i \le b + M(1-y)$$with the $a_i$ coefficients, the $x_i$ variables (discrete, continuous or a mix), $y$ a binary variable and $M$ a "large" coefficient. Think of $M$ as a stunt double for infinity. The notion is that if $y=1$, the constraint is "turned on" and reduces to$$\sum_i a_i x_i \le b.$$If $y=0$, the constraint is "turned off" and reduces to$$\sum_i a_i x_i \le b + M \approx \infty,$$which should be satisfied by any choices for $x$ the solver might make. There are other variations of "big M" constraints, but I'll stick with this one for illustrative purposes.

### The perils of $M$

Choosing a value for $M$ can be a tricky business. Suppose first that we choose $M$ small enough that, when $y=0$ and the constraint should be "off",$$\sum_i a_i x_i^* \gt b + M$$for some solution $x^*$ that

*should*be optimal but now appears infeasible to the solver. The result is an incorrect solution. So let's refer to a value for $M$ as**correct**if it is large enough that no potentially optimal solution violates the constraint when $y=0$. ("Correct" is my term for this. I don't know if there is an "official" term.) Correctness is essential: choose an incorrect (too small) value for $M$ and you risk getting an incorrect solution.Since it is frequently not obvious how large $M$ needs to be in order to be guaranteed correct, people tend to err on the side of caution by choosing really massive values for $M$. That brings with it a different set of problems (ones often ignored in introductory text books). First, branch-and-bound (or branch-and-cut) solvers tend to rely on the continuous relaxation of subproblems (dropping integrality constraints but keeping the functional constraints). Large values of $M$ make for weak relaxations (producing very loose bounds).

Suppose, for instance, that we are solving a model that both builds roads and routes traffic along those roads. $y$ represents the decision to build a particular road ($y=1$) or not ($y=0$). If we build the road, a cost $c$ is incurred (represented by the term $cy$ in the objective function) and we can send unlimited traffic along it; if not, there is no cost but no traffic is committed. In our constraint, the left side is traffic on the road and $b=0$, so traffic there can either be up to $M$ (if built) or 0 (if not built). Now suppose, for example, that the user chooses $M=10^9$ and the solver, in a continuous relaxation of some subproblem, sets $y=10^{-6}$. The solution to the relaxed problem pays only one millionth of $c$ for the privilege of allowing 1,000 units of traffic on this route, which basically allows it to "cheat" (and perhaps grossly underestimates the cost of any actual solution).

A related problem is limited arithmetic precision. Since double-precision floating-point arithmetic (the way the computer does arithmetic with real numbers) has limited precision, and is thus susceptible to rounding errors, solvers have to establish standards for "close enough to feasible" and "close enough to integer". Continuing the previous example, it is entirely possible that the solver might look at $y=10^{-6}$, decide that is within integrality tolerance, and treat it as $y=0$, possibly leading to what it thinks is an "optimal" solution with 1,000 units of traffic running over a road that does not exist. Oops.

Finally, note that $M$ is part of the coefficient matrix. Mixing very large coefficients (like $M$) with much smaller coefficients can create numerical instability, leading the solver to spend more time computing linear program pivots and possibly leading to totally erroneous solutions. That's too big a topic to get into here, but I'm pretty sure I've mentioned it elsewhere.

Despite all this, "big M" models are still very common in practice. There is a nice paper by Codato and Fischetti [1] that shows how a form of Bender's decomposition can be used to get rid of the $M$ coefficients. I've used it successfully (for instance, in [2]), but Bender's decomposition is an advanced technique (i.e., not for the faint of heart), and is not always supported by high-level modeling languages.

So, if we are stuck with "big M" models, what can we do to choose values of $M$ that are both correct and and

**tight**(again, my term), meaning small enough that they avoid numerical problems and hopefully produce relaxations with bounds strong enough to do some good?### Not all $M$ are created equal

Most "big M" models have more than one constraint (and frequently many) containing a large coefficient $M$. One of my pet peeves is that authors of text books and journal articles will frequently just us $M$, with no subscripts, everywhere such a coefficient is needed. This, in turn, breeds a tendency for modelers to choose one large value for $M$ and use it everywhere.

Back when manuscripts were produced on typewriters, it was a bit of a pain in the ass to put in subscripts, so I can see how the trend would have started. (Quick tangent: Nate Brixius recently did a blog post with a picture of a typewriter, for those too young to be familiar with them. I'll link it here for the benefit of younger readers ... and also note that the one pictured is much more modern than the one I learned on, which was not electric.) Today, when people routinely use LaTeX to write manuscripts, there's not much excuse for omitting one measly subscript here and there. Anyway, my point is that it is usually better to use a different, customized value of $M$ for each constraint that needs one.

### Customized how?

In some cases, model context will provide you an obvious choice for $M$. For example, suppose you are selecting warehouse locations and planning shipments from warehouses to customers. A typical "big M" constraint will look like the following (slightly different from our previous constraint but clearly related:$$\sum_j x_{ij} \le M_i y_i.$$Here variable $x_{ij}$ indicates the amount shipped from warehouse $i$ to customer $j$, binary variable $y_i$ is 1 if we use that warehouse and 0 if we do not, and the intent of the constraint is to say that if we do not use (and pay for) a warehouse, we cannot ship anything from it. One obvious choice for $M_i$ is the capacity of the warehouse. A better choice might be the smaller of that capacity or the maximum volume of customers demands that might plausibly be assigned to that warehouse. The latter might be based, for example, on knowing that the company would not ship anything to customers outside a certain distance from the warehouse.

In other cases, there may be some less obvious way to extract suitable (correct and hopefully tight) values for the $M_i$. A while back, I was working on mixed-integer programming models for two group discriminant analysis, and in one paper ([3]) I needed to pick values for $M$. Without going into gory details, I was able to normalize the coefficients of the (linear) discriminant function under construction and then compute valid coefficients $M_i$ by looking at the euclidean distances between pairs of observations. I don't claim that the $M_i$ I came up with were the tightest possible, but they were correct and they produced faster solution of the models than what I got with some arbitrarily large values I initially tried.

Finally, you can actually solve subproblems to get correct and (hopefully) tight $M$ values. Whether this is feasible depends on how many you need and how large and scary your model is. Going back to my first example, the trick would work something like this. Relax all integrality constraints. Drop the constraint in question. If you have any other "big M" constraints for which you have not yet computed tight values of $M$, pick something safely large for those coefficients, trying to avoid numerical problems but not worrying about tightness. Now switch the objective to maximizing $\sum_i a_i x_i - b$. The optimal objective value is your value for $M$. It is clearly correct: any feasible solution to the MIP model is a feasible solution to the LP relaxation, and so the value of $\sum_i a_i x_i - b$ cannot exceed your choice of $M$ (regardless of whether $y$ is 0 or 1). Repeat for each additional constraint containing a "big M" coefficient (switching from minimizing to maximizing if the nature of the constraint warrants it).

The process may be time-consuming, but at least you are solving LPs rather than MIPs. It is also a bit trickier than I made it sound, at least when multiple $M_i$ are involved. You have to guess values for any $M_i$ for which you have not yet solved, and guessing big values for them may result in a looser relaxation than necessary, which in turn may result in an inflated value for the $M_i$ you are currently choosing. You should definitely get correct choices for the $M$ coefficients; it's just the tightness that is in question. Even so, the values you get for the $M_i$ just might be tight enough to save you more time in the MIP solution than it costs to solve the LPs.

### References

[1] Codato, G. and Fischetti, M.

*Combinatorial Benders' Cuts for Mixed-Integer Linear Programming*. Operations Research,**2006**, Vol. 54(4), pp. 756-766.[2] Bai, L. and Rubin, P. A.

*Combinatorial Benders Cuts for the Minimum Tollbooth Problem*. Operations Research,**2009**, Vol. 57(6), pp. 1510-1522.[3] Rubin, P. A.

*Heuristic Solution Procedures for a Mixed-Integer Programming Discriminant Model*. Managerial and Decision Economics,**1990**, Vol. 11(4), pp. 255-266.
Labels:
integer programming,
modeling,
operations research

## Saturday, August 25, 2018

### Adding Items to a Sequence

A question posed on OR-Exchange in 2017 asked the following: Given a tour of nodes, how does one best add two new nodes while respecting the ordering of the original tour. Specifically, the author began with a tour 0 - 1 - 2 - 4 - 6 - 0 (where node 0 is a depot) and wanted to add new stops 3 and 5 in such away that, in the revised tour, stop 1 still came before stop 2, stop 2 before stop 4, etc.

This problem can arise not just in vehicle routing but in many sorts of sequencing problems (such as scheduling jobs for production). Of course, preserving the original ordering to the extent possible is not always a concern, but it might be if, for instance, the existing stops are customers who have been promised somewhat general time windows for delivery. In any event, we'll just take the question as a given.

The answer I posted on OR-X made the somewhat charitable (and, in hindsight, unwarranted) assumption that the two new stops would be inserted by breaking two previous arcs, rather than consecutively (for instance, ... - 2 - 3 - 5 - 4 - ...). So I'll post an answer without that assumption here. In fact, I'll post three variants, one specific to the case of adding exactly two stops and the other two more general.

First, let me articulate some common elements. I'll denote the set of original nodes by $N_1$, the set of nodes to be added by $N_2$, and their union by $N=N_1 \cup N_2$. All three approaches will involve setting up integer programming models that will look for the most part like familiar routing models. So we will have binary variables $x_{ij}$ that will take the value 1 if $j$ immediately follows $i$ in the new tour. We will have constraints ensuring that every node is entered and exited exactly once:$$\sum_{j\in N} x_{ij} = 1\quad \forall i\in N\\ \sum_{i \in N} x_{ij} = 1 \quad\forall j\in N.$$The objective function will be some linear combination of the variables (sum of distances covered, sum of travel times, ...), which I will not worry about here, since it is no different from any sequencing model.

The first new wrinkle is that we do

The first variant makes explicit use of the fact that we have only two new nodes. We add one subtour elimination constraint, to prevent the new nodes from forming a subtour: $x_{35}+x_{53}\le 1.$ Now consider how many different ways we could insert the two new nodes. First, we could break two links in the original tour, inserting 3 in the void where the first link was and 5 in the void where the second link was. Since the original tour had five links there are $\binom{5}{2}=10$ distinct ways to do this. Similarly, we could break two links but insert 5 first and 3 later. There are again ten ways to do it. Finally, we could break one link and insert either 3 - 5 or 5 - 3 into the void. With five choices of the link to break and two possible orders, we get another ten results, for a grand total of 30 possibly new tours.

With that in mind, consider what happens if node 3 is inserted after original node $i$, breaking the link between $i$ and its original successor $j$. (In our model, this corresponds to $x_{i3}=1$.) If this is a single node insertion, then we should have $j$ follow node 3 ($x_{3j}=1$). If it is a double insertion ($i$ - 3 - 5 - $j$), we should have $x_{35}=x_{5j}=1$. We can capture that logic with a pair of constraints for each original arc:

\[ \left.\begin{aligned}x_{i3}-x_{3j} & \le x_{35}\\ x_{i3}-x_{3j} & \le x_{5j} \end{aligned} \right\} \forall(i,j)\in T. \] We could do the same using node 5 in place of node 3, but it is unnecessary. If node 3 is correctly inserted by itself, say between $i$ and $j$, and node 5 is inserted after original node $h$, then the original successor $k$ of $h$ needs a new predecessor. That predecessor cannot be $h$, nor can it be any other original node (given our reduced set of variables), nor can it be node 3 (which now precedes $j$). The only available predecessor is 5, giving us $h$ - 5 - $k$ as expected.

You might wonder how this accommodates a 5 - 3 insertion, say after node $i$. The original successor $j$ of $i$ needs a new predecessor, and 3 is the only eligible choice, so we're good.

I tested this with a small Java program, and it did in fact find all 30 valid revised tours (and no invalid ones).

Version 2, which can be applied to scenarios with any number of new nodes, involves building a standard sequencing model with subtour elimination constraints. The only novel element is the reduced set of variables (as described above). A blog is no place to explain sequencing models in their full glory, so I'll just assume that you, the poor suffering reader, already know how.

In version 3, we again build a sequencing model with the reduced set of variables, but this time we use the Miller-Tucker-Zemlin method of eliminating subtours rather than adding a gaggle of subtour elimination constraints. The MTZ approach generally results in smaller models (since the number of subtours, and hence the potential number of subtour constraints, grows combinatorially with the number of nodes), but also generally produces weaker relaxations.

The Wikipedia page for the TSP shows the MTZ constraints, although for some reason without labeling them as such. Assume a total of $n$ nodes (with consecutive indices), with node $0$ being the depot. The MTZ approach adds continuous variables $u_i, \,i\in \{1,\dots,n\}$ with bounds $0\le u_i \le n-1$. It also adds the following constraints for all eligible arcs $(i,j)$ with $i\neq 0$:$$u_i - u_j + n x_{ij} \le n-1.$$You can think of the $u_i$ variables as counters. The MTZ constraints say that if we go from any node $i$ (other than the depot) to any node $j$ (including the depot), the count at node $j$ has to be at least one larger than the count at node $i$. These constraints preclude any subtours, since a subtour (one starting and ending any place other than the depot) would result in the count at the first node of the subtour being larger than itself.

As I mentioned, the MTZ formulation has a somewhat weaker LP relaxation than a formulation with explicit subtour elimination constraints, so it is not favored by everyone. In our particular circumstance, however, it has an additional virtue: it gives us a relatively painless way to enforce the order preservation requirement. All we need do is insert constraints of the form$$u_j \ge u_i + 1\quad\forall (i,j)\in T.$$This forces the counts at the original nodes to increase monotonically with the original tour order, without directly impacting the counts at the new nodes.

This problem can arise not just in vehicle routing but in many sorts of sequencing problems (such as scheduling jobs for production). Of course, preserving the original ordering to the extent possible is not always a concern, but it might be if, for instance, the existing stops are customers who have been promised somewhat general time windows for delivery. In any event, we'll just take the question as a given.

The answer I posted on OR-X made the somewhat charitable (and, in hindsight, unwarranted) assumption that the two new stops would be inserted by breaking two previous arcs, rather than consecutively (for instance, ... - 2 - 3 - 5 - 4 - ...). So I'll post an answer without that assumption here. In fact, I'll post three variants, one specific to the case of adding exactly two stops and the other two more general.

First, let me articulate some common elements. I'll denote the set of original nodes by $N_1$, the set of nodes to be added by $N_2$, and their union by $N=N_1 \cup N_2$. All three approaches will involve setting up integer programming models that will look for the most part like familiar routing models. So we will have binary variables $x_{ij}$ that will take the value 1 if $j$ immediately follows $i$ in the new tour. We will have constraints ensuring that every node is entered and exited exactly once:$$\sum_{j\in N} x_{ij} = 1\quad \forall i\in N\\ \sum_{i \in N} x_{ij} = 1 \quad\forall j\in N.$$The objective function will be some linear combination of the variables (sum of distances covered, sum of travel times, ...), which I will not worry about here, since it is no different from any sequencing model.

The first new wrinkle is that we do

*not*define a variable for every pair of nodes. We create $x_{ij}$ only for the following combinations of subscripts: \begin{align*} i & \in N_{2},j\in N_{2},i\neq j\\ i & \in N_{1},j\in N_{2}\\ i & \in N_{2},j\in N_{1}\\ i & \in N_{1},j\in N_{1},(i,j)\in T \end{align*} where $T$ is the original tour. Thus, for example, we would have $x_{24}$ but not $x_{42}$, nor $x_{26}$. The rationale is straightforward: if we add an arc between two original nodes that were not successors on the original tour, we will force an order reversal. For instance, suppose we replace the arc 2 - 4 with, say, 2 - 6. Node 4 now must appear either before node 2 or after node 6, and either way the order has not been preserved.### Version 1

The first variant makes explicit use of the fact that we have only two new nodes. We add one subtour elimination constraint, to prevent the new nodes from forming a subtour: $x_{35}+x_{53}\le 1.$ Now consider how many different ways we could insert the two new nodes. First, we could break two links in the original tour, inserting 3 in the void where the first link was and 5 in the void where the second link was. Since the original tour had five links there are $\binom{5}{2}=10$ distinct ways to do this. Similarly, we could break two links but insert 5 first and 3 later. There are again ten ways to do it. Finally, we could break one link and insert either 3 - 5 or 5 - 3 into the void. With five choices of the link to break and two possible orders, we get another ten results, for a grand total of 30 possibly new tours.

With that in mind, consider what happens if node 3 is inserted after original node $i$, breaking the link between $i$ and its original successor $j$. (In our model, this corresponds to $x_{i3}=1$.) If this is a single node insertion, then we should have $j$ follow node 3 ($x_{3j}=1$). If it is a double insertion ($i$ - 3 - 5 - $j$), we should have $x_{35}=x_{5j}=1$. We can capture that logic with a pair of constraints for each original arc:

\[ \left.\begin{aligned}x_{i3}-x_{3j} & \le x_{35}\\ x_{i3}-x_{3j} & \le x_{5j} \end{aligned} \right\} \forall(i,j)\in T. \] We could do the same using node 5 in place of node 3, but it is unnecessary. If node 3 is correctly inserted by itself, say between $i$ and $j$, and node 5 is inserted after original node $h$, then the original successor $k$ of $h$ needs a new predecessor. That predecessor cannot be $h$, nor can it be any other original node (given our reduced set of variables), nor can it be node 3 (which now precedes $j$). The only available predecessor is 5, giving us $h$ - 5 - $k$ as expected.

You might wonder how this accommodates a 5 - 3 insertion, say after node $i$. The original successor $j$ of $i$ needs a new predecessor, and 3 is the only eligible choice, so we're good.

I tested this with a small Java program, and it did in fact find all 30 valid revised tours (and no invalid ones).

### Version 2

Version 2, which can be applied to scenarios with any number of new nodes, involves building a standard sequencing model with subtour elimination constraints. The only novel element is the reduced set of variables (as described above). A blog is no place to explain sequencing models in their full glory, so I'll just assume that you, the poor suffering reader, already know how.

### Version 3

In version 3, we again build a sequencing model with the reduced set of variables, but this time we use the Miller-Tucker-Zemlin method of eliminating subtours rather than adding a gaggle of subtour elimination constraints. The MTZ approach generally results in smaller models (since the number of subtours, and hence the potential number of subtour constraints, grows combinatorially with the number of nodes), but also generally produces weaker relaxations.

The Wikipedia page for the TSP shows the MTZ constraints, although for some reason without labeling them as such. Assume a total of $n$ nodes (with consecutive indices), with node $0$ being the depot. The MTZ approach adds continuous variables $u_i, \,i\in \{1,\dots,n\}$ with bounds $0\le u_i \le n-1$. It also adds the following constraints for all eligible arcs $(i,j)$ with $i\neq 0$:$$u_i - u_j + n x_{ij} \le n-1.$$You can think of the $u_i$ variables as counters. The MTZ constraints say that if we go from any node $i$ (other than the depot) to any node $j$ (including the depot), the count at node $j$ has to be at least one larger than the count at node $i$. These constraints preclude any subtours, since a subtour (one starting and ending any place other than the depot) would result in the count at the first node of the subtour being larger than itself.

As I mentioned, the MTZ formulation has a somewhat weaker LP relaxation than a formulation with explicit subtour elimination constraints, so it is not favored by everyone. In our particular circumstance, however, it has an additional virtue: it gives us a relatively painless way to enforce the order preservation requirement. All we need do is insert constraints of the form$$u_j \ge u_i + 1\quad\forall (i,j)\in T.$$This forces the counts at the original nodes to increase monotonically with the original tour order, without directly impacting the counts at the new nodes.

Labels:
integer programming,
operations research,
routing,
scheduling

## Tuesday, July 31, 2018

### NP Confusion

I just finished reading a somewhat provocative article on the CIO website, titled "10 reasons to ignore computer science degrees" (when hiring programmers). While I'm not in the business of hiring coders (although I recent was hired as a "student programmer" on a grant -- the Universe has a sense of humor), I find myself suspecting that the author is right about a few points, overstating a few and making a few that are valid for some university CS programs but not for all (or possibly most). At any rate, that's not why I mention it here. What particularly caught my eye was the following paragraph:

It’s rare for a CS major to graduate without getting a healthy dose of NP-completeness and Turing machines, two beautiful areas of theory that would be enjoyable if they didn’t end up creating bad instincts. One biologist asked me to solve a problem in DNA sequence matching and I came back to him with the claim that it was NP-complete, a class of problems that can take a very long time to solve. He didn’t care. He needed to solve it anyway. And it turns out that most NP-complete problems are pretty easy to solve most of the time. There are just a few pathological instances that gum up our algorithms. But theoreticians are obsessed with the thin set that confound the simple algorithms, despite being rarely observed in everyday life.Any of my three regular readers will know that I periodically obsess/carp about NP-fixation, so I'm sympathetic to the tenor of this. At the same time, I have a somewhat mixed reaction to it.

- "... NP-complete, a class of problems that can take a very long time to solve." This is certainly factually correct, and the author thankfully said "can" rather than "will". One thing that concerns me in general, though, is that not everyone grasps that problems in class P, for which polynomial time algorithms are known, can also take a very long time to solve. One reason, of course, is that "polynomial time" means run time is a polynomial function of problem size, and big instances will take longer. Another is that $p(n)=n^{1000}$ is a polynomial ... just not one you want to see as a (possibly attainable) bound for solution time. There's a third factor, though, that I think many people miss: the size of the coefficients (including a constant term, if any) in the polynomial bound for run time. I was recently reading a description of the default sorting algorithm in a common programming language. It might have been the one used in the Java collections package, but don't quote me on that. At any rate, they actually use two different sorting algorithms, one for small collections (I think the size cutoff might have been around 47) and the other for larger collections. The second algorithm has better computational complexity, but each step requires a bit more work and/or the setup is slower, so for small collections the nominally more complex algorithm is actually faster.
- "He didn’t care. He needed to solve it anyway." I love this. It's certainly true that users can ask coders (and modelers) for the impossible, and then get snippy when they can't have it, but I do think that mathematicians (and, apparently, computer scientists) can get a bit too locked into theory. <major digression> As a grad student in math, I took a course or two in ordinary differential equations (ODEs), where I got a taste for the differences between mathematicians and engineers. Hand a mathematician an ODE, and he first tries to prove that it has a solution, then tries to characterize conditions under which the solution is unique, then worries about stability of the solution under changes in initial conditions or small perturbations in the coefficients, etc., ad nauseum. An engineer, faced with the same equation, tries to solve it. If she finds the solution, then obviously one exists. Depending on the nature of the underlying problem, she may or may not care about the existence of multiple solutions, and probably is not too concerned about stability given changes in the parameters (and maybe not concerned about changes in the initial conditions, if she is facing one specific set of initial conditions). If she can't solve the ODE, it won't do her much good to know whether a solution exists or not.</major digression> At any rate, when it comes to optimization problems, I'm a big believer in trying a few things before surrendering (and trying a few optimization approaches before saying "oh, it's NP-hard, we'll have to use my favorite metaheuristic").
- "And it turns out that most NP-complete problems are pretty easy to solve most of the time. There are just a few pathological instances that gum up our algorithms." I find this part a bit misleading. Yes, some NP-complete problems can seem easier to solve than others, but the fundamental issue with NP-completeness or NP-hardness is problem dimension. Small instances of problem X are typically easier to solve than larger instances of problem X (although occasionally the Universe will mess with you on this, just to keep you on your toes). Small instances of problem X are likely easier to solve than large instances of problem Y, even if Y seems the "easier" problem class. Secondarily, the state of algorithm development plays a role. Some NP-complete problem classes have received more study than others, and so we have better tools for them. Bill Cook has a TSP application for the iPhone that can solve what I (a child of the first mainframe era) would consider to be insanely big instances of the traveling salesman problem in minutes. So, bottom line, I don't think a "few pathological instances" are responsible for "gum[ming] up our algorithms". Some people have problem instances of a dimension that is easily, or at least fairly easily, handled. Others may have instances (with genuine real-world application) that are too big for our current hardware and software to handle. That's also true of problems in class P. It's just that nobody ever throws their hands up in the air and quits without trying because a problem belongs to class P.

## Tuesday, July 17, 2018

### Selecting Box Sizes

Someone posted an interesting question about box sizes on Mathematics Stack Exchange. He (well, his girlfriend to be precise) has a set of historical documents that need to be preserved in boxes (apparently using a separate box for each document). He wants to find a solution that minimizes the total surface area of the boxes used, so as to minimize waste. The documents are square (I'll take his word for that) with dimensions given in millimeters.

To start, we can make a few simplifying assumptions.

The first thought that crossed my mind was the possibility of solving the problem by brute force. The author of the original question supplied a data file with document dimensions. There are 1166 documents, with 384 distinct sizes. So the brute force approach would be to look at all $\binom{383}{2} = 73,153$ or $\binom{383}{3} = 9,290,431$ combinations of box sizes (in addition to the largest size), calculate the number of boxes of each size and their combined areas, and then choose the combination with the lowest total. On a decent PC, I'm pretty sure cranking through even 9 million plus combinations will only need a tolerable amount of time.

A slightly more sophisticated approach is to view the problem through the lens of a layered network. There are either three or four layers, representing progressively larger selected box sizes, plus a "layer 0" containing a start node. In the three or four layers other than "layer 0", you put one node for each possible box size, with the following restrictions:

I wanted to confirm that the problem is solvable without special purpose software, so I coded it in Java 8. Although there are plenty of high quality open-source graph packages for Java, I wrote my own node, arc and network classes and my own coding of Dijkstra's shortest path algorithm just to prove a point about not needing external software. You are welcome to grab the source code (including the file of document sizes) from my Git repository if you like.

I ran both the three and four size cases and confirmed that my solutions had the same total surface areas that Erwin got, other than a factor of two (I count both top and bottom; he apparently counts just one of them). How long does it take to solve the problem using Dijkstra's algorithm? Including the time reading the data, the four box version takes about half a second on my decent but not workstation caliber PC. The three box version takes about 0.3 seconds, but of course gives a worse solution (since it is more tightly constrained). This is single-threaded, by the way. Both problem set up and Dijkstra's method are amenable to parallel threading, but that would be overkill given the run times.

So is it wrong to take a fancier modeling approach, along the lines of what Erwin did? Not at all. There are just trade-offs. The modeling approach produces more maintainable code (in the form of mathematical models, using some modeling language like GAMS or AMPL) that are also more easily modified if the use case changes. The brute force and basic network approaches I tried requires no extra software (so no need to pay for it, no need to learn it, ...) and works pretty well for a "one-off" situation where maintainability is not critical.

Mainly, though, I just wanted to make a point that we should not overlook simple (or even brute force) solutions to problems when the problem dimensions are small enough to make them practical ... especially with computers getting more and more powerful each year.

To start, we can make a few simplifying assumptions.

- The height of a box is not given, so we'll assume it is zero, and only consider the top and bottom surfaces of the box. For a box (really, envelope) with side $s$, that makes the total area $2s^2$. If the boxes have uniform height $h$, the area changes to $2s^2 + 4hs$, but the model and algorithm I'll pose are unaffected.
- We'll charitably assume that a document with side $s$ fits in a box with side $s$. In practice, of course, you'd like the box to be at least slightly bigger, so that the document goes in and out with reasonable effort. Again, I'll let the user tweak the size formula while asserting that the model and algorithm work well regardless.

- Only document sizes need be considered as box sizes, i.e. for every selected size at least one document should fit "snugly".
- The number of boxes you need at each selected size equals the number of documents too large to fit in a box of the next smaller selected size but capable of fitting in a box of this size.
- You have to select the largest possible box size (since that is required to store the largest of the documents).

The first thought that crossed my mind was the possibility of solving the problem by brute force. The author of the original question supplied a data file with document dimensions. There are 1166 documents, with 384 distinct sizes. So the brute force approach would be to look at all $\binom{383}{2} = 73,153$ or $\binom{383}{3} = 9,290,431$ combinations of box sizes (in addition to the largest size), calculate the number of boxes of each size and their combined areas, and then choose the combination with the lowest total. On a decent PC, I'm pretty sure cranking through even 9 million plus combinations will only need a tolerable amount of time.

A slightly more sophisticated approach is to view the problem through the lens of a layered network. There are either three or four layers, representing progressively larger selected box sizes, plus a "layer 0" containing a start node. In the three or four layers other than "layer 0", you put one node for each possible box size, with the following restrictions:

- the last layer contains only a single node, representing the largest possible box, since you know you are going to have to choose that size;
- the smallest node in each layer is omitted from the following layer (since layers go in increasing size order); and
- the largest node in each layer is omitted from the preceding layer (for the same reason).

I wanted to confirm that the problem is solvable without special purpose software, so I coded it in Java 8. Although there are plenty of high quality open-source graph packages for Java, I wrote my own node, arc and network classes and my own coding of Dijkstra's shortest path algorithm just to prove a point about not needing external software. You are welcome to grab the source code (including the file of document sizes) from my Git repository if you like.

I ran both the three and four size cases and confirmed that my solutions had the same total surface areas that Erwin got, other than a factor of two (I count both top and bottom; he apparently counts just one of them). How long does it take to solve the problem using Dijkstra's algorithm? Including the time reading the data, the four box version takes about half a second on my decent but not workstation caliber PC. The three box version takes about 0.3 seconds, but of course gives a worse solution (since it is more tightly constrained). This is single-threaded, by the way. Both problem set up and Dijkstra's method are amenable to parallel threading, but that would be overkill given the run times.

So is it wrong to take a fancier modeling approach, along the lines of what Erwin did? Not at all. There are just trade-offs. The modeling approach produces more maintainable code (in the form of mathematical models, using some modeling language like GAMS or AMPL) that are also more easily modified if the use case changes. The brute force and basic network approaches I tried requires no extra software (so no need to pay for it, no need to learn it, ...) and works pretty well for a "one-off" situation where maintainability is not critical.

Mainly, though, I just wanted to make a point that we should not overlook simple (or even brute force) solutions to problems when the problem dimensions are small enough to make them practical ... especially with computers getting more and more powerful each year.

Labels:
integer programming,
Java,
network models,
operations research

## Friday, July 6, 2018

### Mint 19 Upgrade: Adventures #1-3

I use my laptop as the "canary in the coal mine" when it comes to do operating system upgrades, since there's nothing awesomely important on it. So today I tried upgrading from Linux Mint 18.3 to 19.0. Note that I used the upgrade path, rather than downloading the installer, burning it to a bootable disk, then installing from there. In hindsight, that might have been the faster approach. The upgrade took over an hour, and that's before any debugging.

I hit the first of what will no doubt be several adventures when I reinstalled RStudio desktop and discovered it would not run. Despite the installer saying that all dependencies were satisfied, when I tried to run it from a command line I was told that a library file (libGL.so.1) could not be found.

I'll skip over another hour or so of pointless flailing and cut to the chase scene. It turns out that libGL.so.1 actually was installed on my laptop, as part of the libgl1-mesa-glx package. It was hiding in plain sight in /usr/lib/x86_64-linux-gnu/mesa/. Somehow, that folder had not made it onto the system library path. (I have no idea why.) So I ran the command

and that fixed the problem.

Next up, I couldn't find a text editor! Note that LibreOffice was installed, and was the default program to open text (.txt) files. Huh?? Poking around, I found nano, but xed (the default text editor in Mint 18) and gedit (the previous default editor) were not installed (even though xed was present before the upgrade).

Fixing this was at least (to quote a math prof I had in grad school) "tedious but brutally straightforward". In the software manager, I installed xed ... and xreader, also MIA. For whatever reason, the other X-Apps (xviewer, xplayer and pix) were already installed (as they all should have been).

Mint has a utility (mintsources) that lets you manage the sources (repositories, PPAs etc.) that you use. There is an entry for it in the main menu, but clicking that entry failed to launch the source manager. On the other hand, running the command ("pkexec mintsources") from a terminal worked just fine.

I found the original desktop file at /usr/share/applications/mintsources.desktop (owned by root, with read and write permissions but not execute permission). After a bunch of messing around, I edited the menu entry through the menu editor (by right-clicking the menu entry and selecting "Edit properties"), changing "pkexec mintsources" to "gksudo mintsources". That creating another version at ~/.local/share/applications/mintsources.desktop. After right-clicking the main menu button and clicking "Reload plugins", the modified entry worked. I have no idea why that works but "pkexec mintsources" does not, even though it does from a terminal. I tried editing back to "pkexec", just in case the mere act of editing was what did the trick, but no joy there. So I edited back to "gksudo", which seems to be working ... for now ... until the gremlins return from their dinner break.

I'll close for now with a link to "Solutions for 24 bugs in Linux Mint 19".

### The case of the not-so-missing library file

I hit the first of what will no doubt be several adventures when I reinstalled RStudio desktop and discovered it would not run. Despite the installer saying that all dependencies were satisfied, when I tried to run it from a command line I was told that a library file (libGL.so.1) could not be found.

I'll skip over another hour or so of pointless flailing and cut to the chase scene. It turns out that libGL.so.1 actually was installed on my laptop, as part of the libgl1-mesa-glx package. It was hiding in plain sight in /usr/lib/x86_64-linux-gnu/mesa/. Somehow, that folder had not made it onto the system library path. (I have no idea why.) So I ran the command

sudo ldconfig /usr/lib/x86_64-linux-gnu/mesa

and that fixed the problem.

### Editor? We don't need no stinkin' editor

Next up, I couldn't find a text editor! Note that LibreOffice was installed, and was the default program to open text (.txt) files. Huh?? Poking around, I found nano, but xed (the default text editor in Mint 18) and gedit (the previous default editor) were not installed (even though xed was present before the upgrade).

Fixing this was at least (to quote a math prof I had in grad school) "tedious but brutally straightforward". In the software manager, I installed xed ... and xreader, also MIA. For whatever reason, the other X-Apps (xviewer, xplayer and pix) were already installed (as they all should have been).

### The mystery of the launcher that wouldn't launch

Mint has a utility (mintsources) that lets you manage the sources (repositories, PPAs etc.) that you use. There is an entry for it in the main menu, but clicking that entry failed to launch the source manager. On the other hand, running the command ("pkexec mintsources") from a terminal worked just fine.

I found the original desktop file at /usr/share/applications/mintsources.desktop (owned by root, with read and write permissions but not execute permission). After a bunch of messing around, I edited the menu entry through the menu editor (by right-clicking the menu entry and selecting "Edit properties"), changing "pkexec mintsources" to "gksudo mintsources". That creating another version at ~/.local/share/applications/mintsources.desktop. After right-clicking the main menu button and clicking "Reload plugins", the modified entry worked. I have no idea why that works but "pkexec mintsources" does not, even though it does from a terminal. I tried editing back to "pkexec", just in case the mere act of editing was what did the trick, but no joy there. So I edited back to "gksudo", which seems to be working ... for now ... until the gremlins return from their dinner break.

**Update:**No sooner did I publish this than I found another instance of the same problem. The driver manager would not launch from the main menu. I edited "pkexec" to "gksudo" for that one, and again it worked. I guess "pkexec" is somehow incompatible with the Mint menu (at least on my laptop).I'll close for now with a link to "Solutions for 24 bugs in Linux Mint 19".

## Thursday, July 5, 2018

### Firefox Ate My Bookmarks

This morning, I "upgraded" Firefox 60.0.2 to 61.0.0 on my desktop computer (Linux Mint). When I started the new version, it came to life with the correct default tabs and pages, no menu bar (my reference), and with the bookmark tool bar visible ... but completely empty. Toggling the menu option to display it was unproductive. I restored the most recent backup of the bookmarks, but the tool bar remained empty.

So I killed Firefox, started it in safe mode (no change), then killed it again and restarted it normally. This time the bookmark tool bar was populated with the correct bookmarks and folders. (I don't know if passing through safe mode was necessary. Maybe it just needed another restart after the restoration operation.) Unfortunately, my problems were not yet over. Although I had the correct top-level stuff in the bookmark tool bar, the various folders only had about three items each, regardless of how many were supposed to be in each folder (and, trust me, it was typically more than three).

When you go to restore bookmarks in Firefox, it will show you a list of backup files (I think it keeps the fifteen most recent) and how many items each contains. My recent backups were all listed with 18 to 21 items. Fortunately, I also have Firefox (not yet upgraded) on my laptop (running the same version of Linux Mint), with the same bookmarks. On the laptop, recent backups have

The laptop proved my savior. I manually backed up the bookmarks on it to a file, parked that file on Dropbox just in case, copied it to the desktop and manually restored it. For the moment, at least, I have all my bookmarks back.

In case you're reading this because you're in the same boat, here are the steps to do manual backups. Of course, this will only help if you have your bookmarks intact somewhere. If you're thinking of upgrading Firefox but haven't pulled the trigger yet, you might want to make a manual backup for insurance.

Start with the "hamburger menu" (the button three horizontal parallel lines). From there, click the "Library" option, then "Bookmarks", then "Show All Bookmarks" (at the very bottom). That opens up a window titled "Library". Click the "Import and Backup" drop-down menu, then either "Backup" or "Restore" depending on your intent. Backup will give you a typical file saving dialog. Restore will give you a list of your recent backups and an option at the bottom to select a file. Use that option to navigate to a manual backup.

Once again, software saves me from having a productive morning. :-(

By the way, this bug has already been reported: https://bugzilla.mozilla.org/show_bug.cgi?id=1472127.

So I killed Firefox, started it in safe mode (no change), then killed it again and restarted it normally. This time the bookmark tool bar was populated with the correct bookmarks and folders. (I don't know if passing through safe mode was necessary. Maybe it just needed another restart after the restoration operation.) Unfortunately, my problems were not yet over. Although I had the correct top-level stuff in the bookmark tool bar, the various folders only had about three items each, regardless of how many were supposed to be in each folder (and, trust me, it was typically more than three).

When you go to restore bookmarks in Firefox, it will show you a list of backup files (I think it keeps the fifteen most recent) and how many items each contains. My recent backups were all listed with 18 to 21 items. Fortunately, I also have Firefox (not yet upgraded) on my laptop (running the same version of Linux Mint), with the same bookmarks. On the laptop, recent backups have

*444*items. So either the upgrade messed up the backup files or Firefox 61.0.0 has trouble reading backups from version 60. Heck, maybe Firefox 60 screwed up making the automatic backups on my desktop (but, somehow, not on my laptop).The laptop proved my savior. I manually backed up the bookmarks on it to a file, parked that file on Dropbox just in case, copied it to the desktop and manually restored it. For the moment, at least, I have all my bookmarks back.

In case you're reading this because you're in the same boat, here are the steps to do manual backups. Of course, this will only help if you have your bookmarks intact somewhere. If you're thinking of upgrading Firefox but haven't pulled the trigger yet, you might want to make a manual backup for insurance.

Start with the "hamburger menu" (the button three horizontal parallel lines). From there, click the "Library" option, then "Bookmarks", then "Show All Bookmarks" (at the very bottom). That opens up a window titled "Library". Click the "Import and Backup" drop-down menu, then either "Backup" or "Restore" depending on your intent. Backup will give you a typical file saving dialog. Restore will give you a list of your recent backups and an option at the bottom to select a file. Use that option to navigate to a manual backup.

Once again, software saves me from having a productive morning. :-(

By the way, this bug has already been reported: https://bugzilla.mozilla.org/show_bug.cgi?id=1472127.

## Tuesday, July 3, 2018

### Usefulness of Computer Science: An Example

I thought I would follow up on my June 29 post, "Does Computer Science Help with OR?", by giving a quick example of how exposure to fundamentals of computer science recently helped me.

A current research project involves optimization models containing large numbers of what are basically set covering constraints, constraints of the form $$\sum_{i\in S} x_i \ge 1,$$ where the $x_i$ are binary variables and $S$ is some subset of the set of all possible indices. The constraints are generated on the fly (exactly how is irrelevant here).

In some cases, the same constraint may be generated more than once, since portions of the code run in parallel threads. Duplicates need to be weeded out before the constraints are added to the main integer programming model. Also, redundant constraints may be generated. By that, I mean we may have two cover constraints, summing over sets $S_1$ and $S_2$, where $S_1 \subset S_2$. When that happens, the first constraint implies the second one, so the second (weaker) constraint is redundant and should be dropped.

So there comes a "moment of reckoning" where all the constraints generated by all those parallel threads get tossed together, and duplicate or redundant ones need to be weeded out. That turns out to be a rather tedious, time-consuming operation, which brings me to how the constraints are represented. I'm coding in Java, which has various implementations of a Set interface to represent sets. The coding path of least resistance would be to toss the indices for each constraint into some class implementing that interface (I generally gravitate to HashSet). The Set interface defines an equals() method to test for equality and a containsAll() method to test whether another set is a subset of a given set. So this would be pretty straightforward to code.

The catch lies in performance. I have not found it documented anywhere, but I

A while back, I took the excellent (and free) online course "Algorithms, Part 1", offered by a couple of faculty from my alma mater Princeton University. I believe it was Robert Sedgewick who said at one point (and I'm paraphrasing here) that sorting is cheap, so if you have any inkling it might help, do it. The binary variables in my model represent selection or non-selection of a particular type of object, and I assigned a complete ordering to them in my code. By "complete ordering" I mean that, given two objects $i$ and $j$, I can tell (in constant time) which one is "preferable". Again, the details do not matter, nor does the plausibility (or implausibility) of the order I made up. It just matters that things are ordered.

So rather than just dump subscripts into HashSets, I created a custom class that stores them in a TreeSet, a type of Java set that maintains sort order using the ordering I created. The custom class also provides some useful functions. One of those functions is isSubsetOf(), which does pretty much what it sounds like: A.isSubsetOf(B) returns true if set $A$ is a subset of set $B$ and false if not.

In the isSubsetOf() method, I start with what are called iterators for the two sets $A$ and $B.$ Each starts out pointing to the smallest member of its set, "smallest" defined according to the ordering I specified. If the smallest member of $B$ is bigger than the smallest member of $A$, then the first element of $A$ cannot belong to $B$, and we have our answer: $A\not\subseteq B$. If the smallest element of $B$ is smaller than the smallest element of $A$, I iterate through element of $B$ until either I find a match to the smallest element of $A$ or run out of elements of $B$ (in which case, again, $A\not\subseteq B$). Suppose I do find a match. I bump the iterator for $A$ to find the second smallest element of $A$, then iterate through subsequent members of $B$ (picking up where I left off in $B$, which is important) until, again, I get a match or die trying. I keep doing this until I get an answer or run out of elements of $A$. At that point, I know that $A\subseteq B$.

What's the payoff for all this extra work? Since I look at each element of $A$ and each element of $B$ at most once, my isSubstOf() method requires $O(n)$ time, not $O(n^2)$ time. Using a TreeSet means the contents of each set have to be sorted at the time of creation, which is $O(n\log n)$, still better than $O(n^2)$. I actually did code it both ways (HashSet versus my custom class) and timed them on one or two moderately large instances. My way is in fact faster. Without having a bit of exposure to computer science (including the Princeton MOOC), though, it would never have occurred to me that I could speed up what was proving to be a bottleneck in my code.

A current research project involves optimization models containing large numbers of what are basically set covering constraints, constraints of the form $$\sum_{i\in S} x_i \ge 1,$$ where the $x_i$ are binary variables and $S$ is some subset of the set of all possible indices. The constraints are generated on the fly (exactly how is irrelevant here).

In some cases, the same constraint may be generated more than once, since portions of the code run in parallel threads. Duplicates need to be weeded out before the constraints are added to the main integer programming model. Also, redundant constraints may be generated. By that, I mean we may have two cover constraints, summing over sets $S_1$ and $S_2$, where $S_1 \subset S_2$. When that happens, the first constraint implies the second one, so the second (weaker) constraint is redundant and should be dropped.

So there comes a "moment of reckoning" where all the constraints generated by all those parallel threads get tossed together, and duplicate or redundant ones need to be weeded out. That turns out to be a rather tedious, time-consuming operation, which brings me to how the constraints are represented. I'm coding in Java, which has various implementations of a Set interface to represent sets. The coding path of least resistance would be to toss the indices for each constraint into some class implementing that interface (I generally gravitate to HashSet). The Set interface defines an equals() method to test for equality and a containsAll() method to test whether another set is a subset of a given set. So this would be pretty straightforward to code.

The catch lies in performance. I have not found it documented anywhere, but I

*suspect*that adding elements to a HashSet is $O(n)$ while checking subset status or equality is $O(n^2)$, where $n$ is the number of possible objects (indices). The reason I say $O(n^2)$ for the latter two operations is that, in the worst case, I suspect that Java takes each object from one subset and compares it to every object in the other set until it finds a match or runs out of things to which to compare. That means potentially $O(n)$ comparisons for each of $O(n)$ elements of the first set, getting us to $O(n^2)$.A while back, I took the excellent (and free) online course "Algorithms, Part 1", offered by a couple of faculty from my alma mater Princeton University. I believe it was Robert Sedgewick who said at one point (and I'm paraphrasing here) that sorting is cheap, so if you have any inkling it might help, do it. The binary variables in my model represent selection or non-selection of a particular type of object, and I assigned a complete ordering to them in my code. By "complete ordering" I mean that, given two objects $i$ and $j$, I can tell (in constant time) which one is "preferable". Again, the details do not matter, nor does the plausibility (or implausibility) of the order I made up. It just matters that things are ordered.

So rather than just dump subscripts into HashSets, I created a custom class that stores them in a TreeSet, a type of Java set that maintains sort order using the ordering I created. The custom class also provides some useful functions. One of those functions is isSubsetOf(), which does pretty much what it sounds like: A.isSubsetOf(B) returns true if set $A$ is a subset of set $B$ and false if not.

In the isSubsetOf() method, I start with what are called iterators for the two sets $A$ and $B.$ Each starts out pointing to the smallest member of its set, "smallest" defined according to the ordering I specified. If the smallest member of $B$ is bigger than the smallest member of $A$, then the first element of $A$ cannot belong to $B$, and we have our answer: $A\not\subseteq B$. If the smallest element of $B$ is smaller than the smallest element of $A$, I iterate through element of $B$ until either I find a match to the smallest element of $A$ or run out of elements of $B$ (in which case, again, $A\not\subseteq B$). Suppose I do find a match. I bump the iterator for $A$ to find the second smallest element of $A$, then iterate through subsequent members of $B$ (picking up where I left off in $B$, which is important) until, again, I get a match or die trying. I keep doing this until I get an answer or run out of elements of $A$. At that point, I know that $A\subseteq B$.

What's the payoff for all this extra work? Since I look at each element of $A$ and each element of $B$ at most once, my isSubstOf() method requires $O(n)$ time, not $O(n^2)$ time. Using a TreeSet means the contents of each set have to be sorted at the time of creation, which is $O(n\log n)$, still better than $O(n^2)$. I actually did code it both ways (HashSet versus my custom class) and timed them on one or two moderately large instances. My way is in fact faster. Without having a bit of exposure to computer science (including the Princeton MOOC), though, it would never have occurred to me that I could speed up what was proving to be a bottleneck in my code.

Labels:
computing,
Java,
operations research,
programming

## Friday, June 29, 2018

### Does Computer Science Help with OR?

Fair warning: tl/dr.

After reading a blog post yesterday by John D. Cook, "Does computer science help you program?", I decided to throw in my two cents (convert to euros at your own risk) on a related topic: does computer science (which I will extend to including programming) help you as an OR/IE/management science/analytics professional? What follows is a mix of opinion and experience, and hence cannot be tested rigorously. (In other words, it's a blog post.)

The obvious starting point is programming. Do you need to know how to program to work in OR (and related fields -- to cut down on typos, I'm going to use "OR" as an umbrella acronym)? In many (most?) cases, yes! Some people with boring jobs may only have to shove data into a statistics program with a graphical user interface and push buttons, or write optimization models in a high level language (technically a form of coding, but I'll discount that) and summon the optimizer, but many will have to do tasks either too complicated for high level development environments, too quirky, or just not suited to one. Unless you magically find yourself endowed with limitless programming support, sooner or later (most likely sooner) you are likely to need to do some coding. (Note to faculty: even if you have hot and cold running students, don't assume that means adequate programming support. I taught simulation coding to doctoral students in a previous life. One of them wrote a program that gave "spaghetti code" a whole new meaning. It looked like a street map of downtown Rome after an earthquake.)

I've done a ton of coding in my time, so perhaps I'm a bit biased. Still, I cannot recall the last OR project I did (research, application, teaching tool or whatever) that did not involve significant coding.

John D. Cook once did a blog post (I don't have the link at hand) about how many languages an analyst or applied mathematician needed to know. I forget the details, but the gist was "one for every type of programming task you do". So here's my toolkit:

I went a long time knowing hardly anything about this, because I was using earlier computing devices where parallelism was out of the question. Today, though, it is pretty easy to find yourself tackling problems where parallel threads or parallel processes are hard to avoid. This includes, but is not limited to, writing programs with a graphical user interface, where doing heavy number crunching in the main thread will freeze the GUI and seriously frustrate the user. I just finished (knock on virtual wood) a Java program for a research project. During the late stages, while tweaking some code related to a heuristic that runs parallel threads and also uses CPLEX, I managed to (a) fill up main memory once (resulting in disk thrashing and essentially paralyzing not only the program interface but the operating system interface ... meaning I couldn't even kill the program) and (b) crash the Java virtual machine (three times, actually). So, trust me, understanding parallel processing can really be important.

This is what John Cook was after in his informal poll. Here, my answer is that some parts are very useful, others pretty much not useful at all, and maybe some stuff in between.

As a graduate student (in math, on the quarter system), I took a three course sequence required for junior year computer science majors. One course concentrated on how to write operating systems. It's perhaps useful to have a basic knowledge of what an operating system does, but I'm pretty sure an OR person can get that from reading computer magazines or something, without taking a course in it. I've totally forgotten what another one of the courses covered, which suggests that maybe it was not crucial to my professional life.

The third course focused on data structures, and while much of what I learned has probably become obsolete (does anyone still use circular buffers?), it's useful both to know the basics and to understand some of the concerns related to different data structures. More than once, while hacking Java code, I've had to give some thought to whether I would be doing a lot of "give me element 7" (which favors array-like structures) versus "give me the smallest element" (favoring tree-like structures) versus "the same thing is going to get added a zillion times, and you only need to keep one copy" (set interfaces, built over who knows what kind of structure).

Another computer science topic you need to know is computational complexity, for a variety of reasons. The first is that, at least if you work in optimization, you will be bombarded by statements about how this problem or that is NP-ridiculous, or this algorithm or that is NP-annoying. (The latter is total b.s. -- NP-hardness is a property of the problem, not the algorithm. Nonetheless, I occasionally see statements like that.) It's important to know both what NP-hardness is (a sign that medium to large problem instances might get pretty annoying) and what it is not (a sign of the apocalypse, or an excuse for skipping the attempt to get an exact solution and immediately dragging out your favorite metaheuristic). You should also understand what a proof of NP-hardness entails, which is more than just saying "it's an integer program".

Beyond NP silliness, though, you need to understand what $O(\dots)$ complexity means, and in particular the fact that an $O(n^2)$ algorithm is slower than an $O(n \log(n))$ algorithm for large $n$, but possibly faster for small $n$. This can help in choosing alternative ways to do computational tasks in your own code.

This may turn up in a numerical analysis class in a mathematics program, or in a computer science course, but either way you

I may have more to say on this at a later date, but right now my computer is running out of electrons, so I'll quit here.

After reading a blog post yesterday by John D. Cook, "Does computer science help you program?", I decided to throw in my two cents (convert to euros at your own risk) on a related topic: does computer science (which I will extend to including programming) help you as an OR/IE/management science/analytics professional? What follows is a mix of opinion and experience, and hence cannot be tested rigorously. (In other words, it's a blog post.)

### Programming

The obvious starting point is programming. Do you need to know how to program to work in OR (and related fields -- to cut down on typos, I'm going to use "OR" as an umbrella acronym)? In many (most?) cases, yes! Some people with boring jobs may only have to shove data into a statistics program with a graphical user interface and push buttons, or write optimization models in a high level language (technically a form of coding, but I'll discount that) and summon the optimizer, but many will have to do tasks either too complicated for high level development environments, too quirky, or just not suited to one. Unless you magically find yourself endowed with limitless programming support, sooner or later (most likely sooner) you are likely to need to do some coding. (Note to faculty: even if you have hot and cold running students, don't assume that means adequate programming support. I taught simulation coding to doctoral students in a previous life. One of them wrote a program that gave "spaghetti code" a whole new meaning. It looked like a street map of downtown Rome after an earthquake.)

I've done a ton of coding in my time, so perhaps I'm a bit biased. Still, I cannot recall the last OR project I did (research, application, teaching tool or whatever) that did not involve significant coding.

#### Languages

John D. Cook once did a blog post (I don't have the link at hand) about how many languages an analyst or applied mathematician needed to know. I forget the details, but the gist was "one for every type of programming task you do". So here's my toolkit:

- one scripting language, for quick or one-off tasks (R for me; others may prefer Python or whatever);
- one language suited for data manipulation/analysis/graphic (R again for me);
- one language for "substantial" computational tasks (Java for me, C++ for a lot of people, Julia for some recent adopters, MATLAB or equivalent for some people, ...); and
- one language for dealing with databases (SQL for me, although "SQL" is like saying "Indian" ... there are a lot of "dialects").

#### Parallel computation

I went a long time knowing hardly anything about this, because I was using earlier computing devices where parallelism was out of the question. Today, though, it is pretty easy to find yourself tackling problems where parallel threads or parallel processes are hard to avoid. This includes, but is not limited to, writing programs with a graphical user interface, where doing heavy number crunching in the main thread will freeze the GUI and seriously frustrate the user. I just finished (knock on virtual wood) a Java program for a research project. During the late stages, while tweaking some code related to a heuristic that runs parallel threads and also uses CPLEX, I managed to (a) fill up main memory once (resulting in disk thrashing and essentially paralyzing not only the program interface but the operating system interface ... meaning I couldn't even kill the program) and (b) crash the Java virtual machine (three times, actually). So, trust me, understanding parallel processing can really be important.

### "Theoretical" Computer Science

This is what John Cook was after in his informal poll. Here, my answer is that some parts are very useful, others pretty much not useful at all, and maybe some stuff in between.

#### Data structures

As a graduate student (in math, on the quarter system), I took a three course sequence required for junior year computer science majors. One course concentrated on how to write operating systems. It's perhaps useful to have a basic knowledge of what an operating system does, but I'm pretty sure an OR person can get that from reading computer magazines or something, without taking a course in it. I've totally forgotten what another one of the courses covered, which suggests that maybe it was not crucial to my professional life.

The third course focused on data structures, and while much of what I learned has probably become obsolete (does anyone still use circular buffers?), it's useful both to know the basics and to understand some of the concerns related to different data structures. More than once, while hacking Java code, I've had to give some thought to whether I would be doing a lot of "give me element 7" (which favors array-like structures) versus "give me the smallest element" (favoring tree-like structures) versus "the same thing is going to get added a zillion times, and you only need to keep one copy" (set interfaces, built over who knows what kind of structure).

#### Complexity

Another computer science topic you need to know is computational complexity, for a variety of reasons. The first is that, at least if you work in optimization, you will be bombarded by statements about how this problem or that is NP-ridiculous, or this algorithm or that is NP-annoying. (The latter is total b.s. -- NP-hardness is a property of the problem, not the algorithm. Nonetheless, I occasionally see statements like that.) It's important to know both what NP-hardness is (a sign that medium to large problem instances might get pretty annoying) and what it is not (a sign of the apocalypse, or an excuse for skipping the attempt to get an exact solution and immediately dragging out your favorite metaheuristic). You should also understand what a proof of NP-hardness entails, which is more than just saying "it's an integer program".

Beyond NP silliness, though, you need to understand what $O(\dots)$ complexity means, and in particular the fact that an $O(n^2)$ algorithm is slower than an $O(n \log(n))$ algorithm for large $n$, but possibly faster for small $n$. This can help in choosing alternative ways to do computational tasks in your own code.

#### Finite precision

This may turn up in a numerical analysis class in a mathematics program, or in a computer science course, but either way you

*really*need to understand what rounding and truncation error are, how accurate double-precision floating point arithmetic is, etc. First, this will help you avoid embarrassment by asking questions like "Why does the solver say my 0-1 variable is 0.999999999999, which is clearly an error?". Second, it will give you an appreciation for why doing things with "stiff" matrices can create remarkably goofy results from very straightforward looking problems. That, in turn, will help you understand why "big M" methods are both a boon and a curse in optimization.I may have more to say on this at a later date, but right now my computer is running out of electrons, so I'll quit here.

Labels:
computing,
operations research,
programming

## Tuesday, June 19, 2018

### Callback Cuts That Repeat

The following post is specific to the CPLEX integer programming solver. I have no idea whether it applies to other solvers, or even which other solver have cut callbacks.

Every so often, a user will discover that a callback routine they wrote has "rediscovered" a cut it previously generated. This can be a bit concerning at first, for a few reasons: it seems implausible, and therefore raises concerns of a bug someplace; it represents repeated work, and thus wasted effort; and it suggests at least the possibility of getting stuck in a loop, the programmatic equivalent of "Groundhog Day". (Sorry, couldn't resist.) As it happens, though, repeating the same cut can legitimately happen (though hopefully not too often).

First, I should probably define what I mean by callbacks here (although I'm tempted to assume that if you don't know what I mean, you've already stopped reading). If you want to get your geek on, feel free to wade through the Wikipedia explanation of a callback function. In the context of CPLEX, what I'm referring to is a user-written function that CPLEX will call at certain points in the search process. I will focus on functions called when CPLEX is generating cuts to tighten the bound at a given node of the search tree (a "user cut callback" in their terminology) or when CPLEX thinks it has identified an integer-feasible solution better than the current incumbent (a "lazy constraint callback"). That terminology pertains to versions of CPLEX prior to 12.8, when those would be two separate (now "legacy") callbacks. As of CPLEX version 12.8, there is a single ("generic") type of callback, but generic callbacks continue to be called from multiple "contexts", including those two (solving a node LP, checking a new candidate solution).

The purpose here is to let the user either generate a bound-tightening constraint ("user cut") using particular knowledge of the problem structure, or to vet a candidate solution and, if unsuitable, issue a new constraint ("lazy constraint") that cuts off that solution, again based on problem information not part of the original IP model. The latter is particularly common with decomposition techniques such as Benders decomposition.

So why would the same user cut or lazy constraint be generated more than once (other than due to a bug)? There are at least two, and possibly three, explanations.

A user cut can be either local or global. The difference is whether the cut is valid in the original model ("global") or whether it is contingent on branching decisions that led to the current node ("local"). The user can specify in the callback whether a new cut is global or local. If the cut is specified as local, it will be enforced only at descendants of the current node.

That said, it is possible that a local cut might be valid at more than one node of the search tree, in which case it might be rediscovered when those other nodes are visited. In the worst case, if the cut is actually globally valid but for some reason added with the local flag set, it may be rediscovered quite a few times.

On a system with multiple cores (or multiple processors), using parallel threads can speed CPLEX up. Parallel threads, though, are probably the most common cause for cuts and lazy constraints repeating.

The issue is that threads operate somewhat independently and only synchronize periodically. So suppose that thread A triggers a callback that generates a globally valid user cut or lazy constraint, which is immediately added to the problem A is solving. Thread B, which is working on a somewhat different problem (from a different part of the search tree), is unaware of the new cut/constraint until it reaches a synchronization point (and finds out what A and any other sibling threads have been up to). Prior to that, B might stumble on the same cut/constraint. Since A and B are calling the same user callback function, if the user is tracking what has been generated inside that function, the function will register a repetition. This is normal (and harmless).

This last explanation is one I am not entirely sure about. When cuts and lazy constraints are added, CPLEX stores them internally in some sort of table. I believe that it is possible in some circumstances for a callback function to be called before the table is (fully) checked, in which case the callback might generate a cut or constraint that is already in the table. Since this deals with the internal workings of CPLEX (the secret sauce), I don't know first-hand if this is true or not ... but if it is, it is again slightly wasteful of time but generally harmless.

Of course, none of this precludes the possibility of a bug in the user's code. If, for example, the user reacts to a candidate solution with a lazy constraint that is intended to cut off that solution but does not (due to incorrect formulation or construction), CPLEX will register the user constraint, notice that the solution is still apparently valid, and give the callback another crack at it. (At least that is how it worked with legacy callbacks, and I

Every so often, a user will discover that a callback routine they wrote has "rediscovered" a cut it previously generated. This can be a bit concerning at first, for a few reasons: it seems implausible, and therefore raises concerns of a bug someplace; it represents repeated work, and thus wasted effort; and it suggests at least the possibility of getting stuck in a loop, the programmatic equivalent of "Groundhog Day". (Sorry, couldn't resist.) As it happens, though, repeating the same cut can legitimately happen (though hopefully not too often).

First, I should probably define what I mean by callbacks here (although I'm tempted to assume that if you don't know what I mean, you've already stopped reading). If you want to get your geek on, feel free to wade through the Wikipedia explanation of a callback function. In the context of CPLEX, what I'm referring to is a user-written function that CPLEX will call at certain points in the search process. I will focus on functions called when CPLEX is generating cuts to tighten the bound at a given node of the search tree (a "user cut callback" in their terminology) or when CPLEX thinks it has identified an integer-feasible solution better than the current incumbent (a "lazy constraint callback"). That terminology pertains to versions of CPLEX prior to 12.8, when those would be two separate (now "legacy") callbacks. As of CPLEX version 12.8, there is a single ("generic") type of callback, but generic callbacks continue to be called from multiple "contexts", including those two (solving a node LP, checking a new candidate solution).

The purpose here is to let the user either generate a bound-tightening constraint ("user cut") using particular knowledge of the problem structure, or to vet a candidate solution and, if unsuitable, issue a new constraint ("lazy constraint") that cuts off that solution, again based on problem information not part of the original IP model. The latter is particularly common with decomposition techniques such as Benders decomposition.

So why would the same user cut or lazy constraint be generated more than once (other than due to a bug)? There are at least two, and possibly three, explanations.

### Local versus Global

A user cut can be either local or global. The difference is whether the cut is valid in the original model ("global") or whether it is contingent on branching decisions that led to the current node ("local"). The user can specify in the callback whether a new cut is global or local. If the cut is specified as local, it will be enforced only at descendants of the current node.

That said, it is possible that a local cut might be valid at more than one node of the search tree, in which case it might be rediscovered when those other nodes are visited. In the worst case, if the cut is actually globally valid but for some reason added with the local flag set, it may be rediscovered quite a few times.

### Parallel Threading

On a system with multiple cores (or multiple processors), using parallel threads can speed CPLEX up. Parallel threads, though, are probably the most common cause for cuts and lazy constraints repeating.

The issue is that threads operate somewhat independently and only synchronize periodically. So suppose that thread A triggers a callback that generates a globally valid user cut or lazy constraint, which is immediately added to the problem A is solving. Thread B, which is working on a somewhat different problem (from a different part of the search tree), is unaware of the new cut/constraint until it reaches a synchronization point (and finds out what A and any other sibling threads have been up to). Prior to that, B might stumble on the same cut/constraint. Since A and B are calling the same user callback function, if the user is tracking what has been generated inside that function, the function will register a repetition. This is normal (and harmless).

### Cut Tables

This last explanation is one I am not entirely sure about. When cuts and lazy constraints are added, CPLEX stores them internally in some sort of table. I believe that it is possible in some circumstances for a callback function to be called before the table is (fully) checked, in which case the callback might generate a cut or constraint that is already in the table. Since this deals with the internal workings of CPLEX (the secret sauce), I don't know first-hand if this is true or not ... but if it is, it is again slightly wasteful of time but generally harmless.

### User Error

Of course, none of this precludes the possibility of a bug in the user's code. If, for example, the user reacts to a candidate solution with a lazy constraint that is intended to cut off that solution but does not (due to incorrect formulation or construction), CPLEX will register the user constraint, notice that the solution is still apparently valid, and give the callback another crack at it. (At least that is how it worked with legacy callbacks, and I

*think*that is how it works with generics.) Seeing the same solution, the user callback might generate the same (incorrect) lazy constraints, and off the code goes, chasing its own tail.## Wednesday, May 30, 2018

### Getting Videos to Play in Firefox

I

I will periodically land on a web page with a video at the top. Usually I don't want to play the video, but sometimes I do ... in which case I have mixed results getting the video to play in Firefox. Typically Chrome has no such problem with the page, which makes sense given that the issue in Firefox relates to a couple of browser extensions I use. What made diagnosing the problem tricky was that (a) results were inconsistent (some videos played, some didn't) and (b) the problem resulted from a combination of two extensions, not just one particular one.

They symptom was that, when I clicked the play button on a video, the player box would turn into a black screen with a spinner that would spin indefinitely, until an error message popped up saying that the video player had hit a time limit waiting for the video to load. Again, this happened on some videos but not others. In particular, I never had a problem with a YouTube video. It was at least consistent in that a video that triggered the error would always trigger the error, regardless of page reloads etc.

The two Firefox extensions involved are HTTPS Everywhere (which tries to force page contents to load using the more secure HTTPS protocol than the ordinary HTTP protocol) and Disable HTML5 Autoplay (which prevents videos from automatically starting to play). The first extension is a security enhancement. (For an opinion piece on why HTTPS is important, read "Why HTTPS Matters".) Unfortunately, many web sites either do not deploy HTTPS for some content or screw up their site configuration. Regarding the second extension, I find video autoplay to be rather annoying, since it frequently involves ads or other videos that I have no interest in, and forces me to play whack-a-mole to shut them the bleep up.

I'm loathe to give up either extension, and fortunately I don't have to. What works for me is a combination of two tweaks. On a site where I get problem videos (time.com is the main source, in my case), I click the toolbar button for the autoplay extension, leave "Disable Autoplay" selected, but deselect "Disable Preloading". That only needs to be done once per site. With a page open containing a problem video, I then disable HTTPS Everywhere (again, by clicking its toolbar button and deselecting the first option). That should automatically cause the page to reload, and the video will play properly. After I'm done watching, I just reenable HTTPS Everywhere. This part has to be repeated for each page containing a video that will not load via HTTPS, but it's a price I'm willing to pay to preserve security.

*think*I've solved a lingering problem I've had playing certain videos in Firefox, and I'm posting the details here mostly so I don't forget them and partly in case anybody else is tripping over this.I will periodically land on a web page with a video at the top. Usually I don't want to play the video, but sometimes I do ... in which case I have mixed results getting the video to play in Firefox. Typically Chrome has no such problem with the page, which makes sense given that the issue in Firefox relates to a couple of browser extensions I use. What made diagnosing the problem tricky was that (a) results were inconsistent (some videos played, some didn't) and (b) the problem resulted from a combination of two extensions, not just one particular one.

They symptom was that, when I clicked the play button on a video, the player box would turn into a black screen with a spinner that would spin indefinitely, until an error message popped up saying that the video player had hit a time limit waiting for the video to load. Again, this happened on some videos but not others. In particular, I never had a problem with a YouTube video. It was at least consistent in that a video that triggered the error would always trigger the error, regardless of page reloads etc.

The two Firefox extensions involved are HTTPS Everywhere (which tries to force page contents to load using the more secure HTTPS protocol than the ordinary HTTP protocol) and Disable HTML5 Autoplay (which prevents videos from automatically starting to play). The first extension is a security enhancement. (For an opinion piece on why HTTPS is important, read "Why HTTPS Matters".) Unfortunately, many web sites either do not deploy HTTPS for some content or screw up their site configuration. Regarding the second extension, I find video autoplay to be rather annoying, since it frequently involves ads or other videos that I have no interest in, and forces me to play whack-a-mole to shut them the bleep up.

I'm loathe to give up either extension, and fortunately I don't have to. What works for me is a combination of two tweaks. On a site where I get problem videos (time.com is the main source, in my case), I click the toolbar button for the autoplay extension, leave "Disable Autoplay" selected, but deselect "Disable Preloading". That only needs to be done once per site. With a page open containing a problem video, I then disable HTTPS Everywhere (again, by clicking its toolbar button and deselecting the first option). That should automatically cause the page to reload, and the video will play properly. After I'm done watching, I just reenable HTTPS Everywhere. This part has to be repeated for each page containing a video that will not load via HTTPS, but it's a price I'm willing to pay to preserve security.

## Tuesday, May 15, 2018

### Grouping Rows of a Matrix

I spent a large chunk of yesterday afternoon doing something I thought would be simple (relatively speaking) in LaTeX. I wanted to group rows of a matrix (actually, in my case, a vector) with right braces, and label the groups. An example of what I wanted is in the image below.

This seems to me to be a fairly common thing to do, and LaTeX has been around over 35 years (TeX even longer), so by now it must be pretty easy. Right? Um, not so much. I wore out Google looking for packages that would do this. Curiously, it's easy to put braces

Fortunately, something pointed me in the direction of the PGF/TiKZ package, which I've used in the past for doing drawings. It's an incredible tool in terms of both what it can do and the outstanding quality of its manual. Because it does so many things, I've never really gotten to know all its capabilities, and in particular its ability to do matrices in a picture environment.

Here is the code to do my illustration. You need to load the TiKZ package and two of its libraries in your document preamble, as follows:

The code for the drawing is:

The name of the matrix ("vec") is arbitrary. The amplitude for the brace (10pt) and the offsets (10pt and 20pt) are matters of taste.

If you happen to know a faster way of doing this, please do share in a comment.

*over*and*under*things:- $\overbrace{[x_1,\dots,x_n]}$ [\overbrace{[x_1,\dots,x_n]}];
- $\underbrace{[x_1,\dots,x_n]}$ [\underbrace{[x_1,\dots,x_n]}].

Fortunately, something pointed me in the direction of the PGF/TiKZ package, which I've used in the past for doing drawings. It's an incredible tool in terms of both what it can do and the outstanding quality of its manual. Because it does so many things, I've never really gotten to know all its capabilities, and in particular its ability to do matrices in a picture environment.

Here is the code to do my illustration. You need to load the TiKZ package and two of its libraries in your document preamble, as follows:

\usepackage{tikz} \usetikzlibrary{matrix, decorations.pathreplacing}

The code for the drawing is:

\begin{tikzpicture} \matrix (vec) [matrix of math nodes, left delimiter = {[}, right delimiter = {]}] { f_1 \\ \vdots \\ f_{a} \\ f_{a + 1} \\ \vdots \\ f_{b} \\ f_{b + 1} \\ \vdots \\ f_{c} \\ }; \node (a) at (vec-1-1.north) [right=20pt]{}; \node (b) at (vec-3-1.south) [right=20pt]{}; \node (c) at (vec-4-1.north) [right=20pt]{}; \node (d) at (vec-6-1.south) [right=20pt]{}; \node (e) at (vec-7-1.north) [right=20pt]{}; \node (f) at (vec-9-1.south) [right=20pt]{}; \draw [decorate, decoration={brace, amplitude=10pt}] (a) -- (b) node[midway, right=10pt] {\footnotesize something}; \draw [decorate, decoration={brace, amplitude=10pt}] (c) -- (d) node[midway, right=10pt] {\footnotesize something else}; \draw [decorate, decoration={brace, amplitude=10pt}] (e) -- (f) node[midway, right=10pt] {\footnotesize something silly}; \end{tikzpicture}

The name of the matrix ("vec") is arbitrary. The amplitude for the brace (10pt) and the offsets (10pt and 20pt) are matters of taste.

If you happen to know a faster way of doing this, please do share in a comment.

## Friday, April 27, 2018

### Big M and Integrality Tolerance

A change I made to an answer I posted on OR-Exchange, based on a comment from a well-informed user of OR-X, might be worth repeating here on the blog. It has to do with issues that can occur when using "big M" type integer programming models, a topic I've covered here before.

As I mentioned in that previous post, part of the problem in "big M" formulations stems from the inevitable rounding error in any non-integer computations done on a computer. A particular manifestation of rounding error (regardless of whether there are "big M" coefficients in the model or not) is that the double precision value assigned to integer variables in a solution will not necessarily be integers. With surprising frequency, I see users of MIP software demanding to know why the answer they got for their integer variable was 0.9999999999975 or 1.000000000032 rather than exactly 1. The answer has two parts: (a) rounding error is pretty much inevitable; and (b) the software designers accepted that reality and decreed that anything "close enough" to the nearest integer counts as being an integer for purposes of deciding if a solution is feasible. (Why the software prints all those decimal places rather than rounding for you is a separate question that I will not address.)

So the solver generally has a parameter that gives an "integrality tolerance", just as it has a (typically separate) parameter for how close the expression in a constraint has to be to the allowed value(s) to be considered feasible (again, a nod to rounding error). In CPLEX, the name of the integrality tolerance parameter is some variant of "MIP.Tolerances.Integrality" (in earlier versions, the much more compact "EpInt"), and its default value (as of this writing) is 1.0E-5.

So now I'll connect that to "big M". One of the more common uses of "big M" is to capture a logical constraint of the form "if condition is true then expression is limited". For instance, you might want to build into the model that if a warehouse is not open (the condition) then shipments from it must equal (or cannot exceed) zero. Algebraically, this frequently appears as $$f(x) \le My$$where $f(x)$ is some (hopefully linear) expression involving variable $x$ and $y$ is a binary variable (with 1 meaning the constraint is relaxed and 0 meaning it is enforced). In a world of exact arithmetic, and with $M$ chosen large enough, $y=1$ means the value of $f(x)$ is essentially unbounded above, while $y=0$ means $f(x)\le 0$.

Here the issue of integrality tolerance sneaks in. Suppose that we choose some really large value of $M$, say $M=1E+10$, and that the solver decides to accept a solution where $y=1.0E-6$ (which is within CPLEX's default integrality tolerance). From the solver's perspective, $y=0$. Logically, that should mean $f(x)\le 0$, but given the rounding error in $y$ and the large value of $M$ what you actually get is $f(x)\le 10,000$. So, borrowing from my earlier example, I've got a closed warehouse shipping 10,000 units of whatever. Oops.

A common reaction to this (and by "common" I mean I've seen it multiple times on help forums) is to say "I'll set the integrality tolerance to zero". Good luck with that. First, it's not guaranteed the software will let you. (CPLEX will.) Second, if it does let you do it, you might get a suboptimal solution, or be told your perfectly feasible problem is actually infeasible, because the software couldn't get the true optimal solution (or perhaps any solution) to have zero rounding error in all the integer variables.

If you run into incorrect solutions in a "big M" model, some combination of tightening the integrality tolerance (but not all the way to zero) and ratcheting down the size of $M$ may fix things ... but, as with all aspects of MIP models, there are no guarantees.

As I mentioned in that previous post, part of the problem in "big M" formulations stems from the inevitable rounding error in any non-integer computations done on a computer. A particular manifestation of rounding error (regardless of whether there are "big M" coefficients in the model or not) is that the double precision value assigned to integer variables in a solution will not necessarily be integers. With surprising frequency, I see users of MIP software demanding to know why the answer they got for their integer variable was 0.9999999999975 or 1.000000000032 rather than exactly 1. The answer has two parts: (a) rounding error is pretty much inevitable; and (b) the software designers accepted that reality and decreed that anything "close enough" to the nearest integer counts as being an integer for purposes of deciding if a solution is feasible. (Why the software prints all those decimal places rather than rounding for you is a separate question that I will not address.)

So the solver generally has a parameter that gives an "integrality tolerance", just as it has a (typically separate) parameter for how close the expression in a constraint has to be to the allowed value(s) to be considered feasible (again, a nod to rounding error). In CPLEX, the name of the integrality tolerance parameter is some variant of "MIP.Tolerances.Integrality" (in earlier versions, the much more compact "EpInt"), and its default value (as of this writing) is 1.0E-5.

So now I'll connect that to "big M". One of the more common uses of "big M" is to capture a logical constraint of the form "if condition is true then expression is limited". For instance, you might want to build into the model that if a warehouse is not open (the condition) then shipments from it must equal (or cannot exceed) zero. Algebraically, this frequently appears as $$f(x) \le My$$where $f(x)$ is some (hopefully linear) expression involving variable $x$ and $y$ is a binary variable (with 1 meaning the constraint is relaxed and 0 meaning it is enforced). In a world of exact arithmetic, and with $M$ chosen large enough, $y=1$ means the value of $f(x)$ is essentially unbounded above, while $y=0$ means $f(x)\le 0$.

Here the issue of integrality tolerance sneaks in. Suppose that we choose some really large value of $M$, say $M=1E+10$, and that the solver decides to accept a solution where $y=1.0E-6$ (which is within CPLEX's default integrality tolerance). From the solver's perspective, $y=0$. Logically, that should mean $f(x)\le 0$, but given the rounding error in $y$ and the large value of $M$ what you actually get is $f(x)\le 10,000$. So, borrowing from my earlier example, I've got a closed warehouse shipping 10,000 units of whatever. Oops.

A common reaction to this (and by "common" I mean I've seen it multiple times on help forums) is to say "I'll set the integrality tolerance to zero". Good luck with that. First, it's not guaranteed the software will let you. (CPLEX will.) Second, if it does let you do it, you might get a suboptimal solution, or be told your perfectly feasible problem is actually infeasible, because the software couldn't get the true optimal solution (or perhaps any solution) to have zero rounding error in all the integer variables.

If you run into incorrect solutions in a "big M" model, some combination of tightening the integrality tolerance (but not all the way to zero) and ratcheting down the size of $M$ may fix things ... but, as with all aspects of MIP models, there are no guarantees.

## Sunday, March 11, 2018

### Piecewise Linear Approximations in MIP Models

In the past, I've written about piecewise linear approximations of functions of a single variable. (There are too many posts to list here. Just type "piecewise linear" in the blog search box if you want to find them.) Handling piecewise linear approximations of multivariable functions is a bit more intimidating. I'll illustrate one approach here.

To set the stage, assume an optimization problem involving variables $x\in\Re^{n}$, and possibly some other variables $y\in\Re^{m}$. Some or all of the $x$ and $y$ variables may be restricted to integer values. Also assume that the objective function and constraints are all linear, with the exception of one nonlinear function of the $x$ variables that may appear in one or more constraints and/or the objective function. I'll denote that function $f:\Re^{n}\rightarrow\Re$. With that, I'll write our hypothetical problem as

\begin{align*}

\text{min } & g(x,y,z)\\

\text{s.t. } & z=f(x)\\

& (x,y,z)\in D

\end{align*}where $z$ is a new variable that captures the value of $f$, the domain $D$ incorporates all other functional constraints, bounds and integrality restrictions, and the objective function $g$ is linear. So if it were not for the constraint $z=f(x)$, this would be a mixed integer linear program.

I'm also going to assume that we know

I'll use the term "mesh points" for the points$$\prod_{i=1}^{n}\left\{ a_{i}^{(1)},\dots,a_{i}^{(N_{i})}\right\}$$formed by combinations of grid values, and $a^{(j)}$ will denote a generic mesh point $(a_{1}^{(j_{1})},\dots,a_{n}^{(j_{n})})$. The superscript $j$ for $w$ is a vector $(j_{1},\dots,j_{n})$ whose domain I will denote $J$. Now we can get down to the piecewise linear approximation. We will write $x$ as a convex combination of the mesh points, and approximate $f(x)$ with the corresponding convex combination of the function values at the mesh points. Just to be clear, in this formulation $z$ approximates, but no longer equals, $f(x)$. To do this, we will introduce new (continuous) variables $w^{(j)}\in[0,1]$ for each $j\in J$. There are $N_{1}\times\cdots\times N_{n}$ mesh points, and so an identical number of $w$ variables. The $w$ variables are weights assigned to the mesh points. This leads to the following additional constraints:

\begin{align*}

\sum_{j\in J} & w^{(j)}=1\\

\sum_{j\in J}a^{(j)}w^{(j)} & =x\\

\sum_{j\in J} & f(a^{(j)})w^{(j)}=z.

\end{align*}

There's still one more wrinkle with which to contend. Other than possibly at extreme points, there will be more than one convex combination of mesh points producing the same $x$ vector, and they will not all produce the same approximation $z$ of $f(x)$. Consider an example in which $n=2$, $R=[2,5]\times[1,3]$, and $f(x)=x^{\prime}x=\left\Vert x\right\Vert ^{2}$. I'll use integer-valued mesh points. Assume that the optimal solution requires that $x=(3.5,2.2)$, in which case $f(x)=17.09$. Figure 1 illustrates the situation, with the values of $f()$ at the mesh points shown in red. (Click any figure to get a better resolution version in a separate browser tab.)

Assume first that the solver prefers smaller values of $f(x)$. Figure 2 shows the weights $w$ that minimize $z$, the estimate of $f(x)$, for our given choice of $x$. The interpolated value of $z$ is $17.5$, which is moderately close to the correct value $17.09$. The three mesh points closest to $x$ receive weights 0.3, 0.5 and 0.2. Figure 3 shows an alternative solution with the same value of $z$, using a different combination of adjacent corners.

Now assume instead that the solver prefers larger values of $f(x)$. The solution that maximizes $z$ is shown in Figure 4. It uses three corners of R, none of which are adjacent to $x,$ with weights 0.5, 0.4 and 0.1 Although this produces the correct value of $x$, the interpolated value of $z$ is $20.3$, which grossly overstates the correct value $17.09$.

To keep the approximation as accurate as possible, we should force the solver to use mesh points adjacent to the actual solution $x$. We can do that by introducing a new binary variable $v_{i}^{(j)}\in\left\{ 0,1\right\} $ for each grid value $a_{i}^{(j)}$. Variable $v_{i}^{(j)}$ will signal whether $a_{i}^{(j)}$ is the $i$-th coordinate of the "lower left" (more generally, closest to the origin) corner of the mesh hyperrectangle containing $x$. Since we want to select a single hyperrectangle to contain $x$, we add constraints requiring exactly one choice for each coordinate of the lower left corner:

\[

\sum_{j=1}^{N_{i}}v_{i}^{(j)}=1\quad i=1,\dots,n.

\]Now we just need to add constraints forcing $w^{(j)}=0$ unless mesh point $j$ is one of the corners of the chosen hyperrectangle. Observe that, along any dimension $i$, the $i$-th component of any corner of the correct hyperrectangle will either be the chosen grid value or the next consecutive grid value. For instance, the rectangle containing $x$ in Figure 1 has lower left corner $(a_{1}^{(2)},a_{2}^{(2)})=(3,2)$. That means $v_{1}^{(2)}=1=v_{2}^{(2)}$. The corners of the rectangle have either $3=a_{1}^{(2)}$ or $4=a_{1}^{(3)}$ for their $x_{1}$ coordinate and either $2=a_{2}^{(2)}$ or $3=a_{2}^{(3)}$ for their $x_{2}$ coordinate.

So for $w^{(j)}$ to be nonzero, we need either $v_{i}^{(j_{i})}=1$ or $v_{i}^{(j_{i}-1)}=1$. This leads us to the constraints \[

w^{(j)}\le v_{i}^{(j_{i})}+v_{i}^{(j_{i-1})}

\]for all indices $j$ of mesh points and for all $i=1,\dots,n$, with the understanding that $v_{i}^{(0)}=0$ (since there is no grid point prior to the first one in any direction).

With those extra constraints, the solutions to our little example when we want $z$ small are unchanged, since they already obey the additional constraints. When we want $z$ large, however, the solution in Figure 4 is now infeasible. All three positive weights violate the new constraints. For instance, $w^{(1,3)}=0.5$ (the weight applied to the mesh point formed by the first grid value of $x_{1}$ and the third grid value of $x_{2}$), but $v_{1}^{(1)}+v_{1}^{(0)}=0$. The solutions that maximize $z$ end up being the same ones that minimize $z$ (those shown in Figures 2 and 3).

What remains is to take stock of how much the model has inflated. Let $$P=N_{1}\times\cdots\times N_{n}$$and$$S=N_{1}+\cdots+N_{n}.$$ We first added $P$ continuous variables ($w$) and $n+2$ constraints involving them. Then we added $S$ binary variables ($v$), $n$ constraints involving just them, and $nP$ constraints tying them to the $w$ variables. That's a grand total of $P$ continuous variables, $S$ binary variables and $2n+2+nP$ constraints. Note that $n\ll S\ll P<nP$. Also note that, in general, continuous variables are computationally cheaper than integer variables. As for the gaggle of extra constraints, modern solvers have ways to deal with some constraints "lazily", which mitigates the load to some extent.

An alternative approach would be to partion $R$ into nonoverlapping polygons (not necessarily hyperrectangles), assign weight variables $w$ to the corners of those polygons as above, and assign a binary variable to each polygon indicating whether it was selected. That would increase the number of binary variables from $S$ to $P$ (where $P$ would be the number of polygons) and decrease the number of added constraints from $2n+2+nP$ to $n+3+P.$ (The $n$ constraints requiring sums of binary variables to be 1 becomes a single constraint summing all the new binary variables. The $nP$ constraints tying the $w$ and $v$ variables together become $P$ constraints, one for each $w$ variable.) This approach is more flexible in terms of concentrating polygons where the curvature of $f$ is greatest, and it significantly reduces the number of constraints; but it significantly increases the number of binary variables, albeit with all of them tied into a single type 1 special ordered set. So it's hard for me to say which is better in general.

To set the stage, assume an optimization problem involving variables $x\in\Re^{n}$, and possibly some other variables $y\in\Re^{m}$. Some or all of the $x$ and $y$ variables may be restricted to integer values. Also assume that the objective function and constraints are all linear, with the exception of one nonlinear function of the $x$ variables that may appear in one or more constraints and/or the objective function. I'll denote that function $f:\Re^{n}\rightarrow\Re$. With that, I'll write our hypothetical problem as

\begin{align*}

\text{min } & g(x,y,z)\\

\text{s.t. } & z=f(x)\\

& (x,y,z)\in D

\end{align*}where $z$ is a new variable that captures the value of $f$, the domain $D$ incorporates all other functional constraints, bounds and integrality restrictions, and the objective function $g$ is linear. So if it were not for the constraint $z=f(x)$, this would be a mixed integer linear program.

I'm also going to assume that we know

*a priori*some hyperrectangle $R\subset\Re^{n}$ containing all feasible values of $x$, say $R=\prod_{i=1}^{n}[L_{i},U_{i}]$. Brace yourself, because the notation is about to get a bit messy. We will create a mesh of discrete points at which $f$ will be evaluated. First, we specify a sequence of what I will call "grid values" for each individual component of $x$. I'll denote the $j$-th grid value for $x_{i}$ by $a_{i}^{(j)}$, where $$L_{i}=a_{i}^{(1)}<a_{i}^{(2)}<\cdots<a_{i}^{(N_{i})}=U_{i}.$$The number of grid values need not be the same for each variable (hence the subscript on $N_{i}$), and the spacing need not be equal. It probably makes sense to sample $f()$ more frequently in regions where its curvature is greater, and less frequently in regions where it is fairly flat.I'll use the term "mesh points" for the points$$\prod_{i=1}^{n}\left\{ a_{i}^{(1)},\dots,a_{i}^{(N_{i})}\right\}$$formed by combinations of grid values, and $a^{(j)}$ will denote a generic mesh point $(a_{1}^{(j_{1})},\dots,a_{n}^{(j_{n})})$. The superscript $j$ for $w$ is a vector $(j_{1},\dots,j_{n})$ whose domain I will denote $J$. Now we can get down to the piecewise linear approximation. We will write $x$ as a convex combination of the mesh points, and approximate $f(x)$ with the corresponding convex combination of the function values at the mesh points. Just to be clear, in this formulation $z$ approximates, but no longer equals, $f(x)$. To do this, we will introduce new (continuous) variables $w^{(j)}\in[0,1]$ for each $j\in J$. There are $N_{1}\times\cdots\times N_{n}$ mesh points, and so an identical number of $w$ variables. The $w$ variables are weights assigned to the mesh points. This leads to the following additional constraints:

\begin{align*}

\sum_{j\in J} & w^{(j)}=1\\

\sum_{j\in J}a^{(j)}w^{(j)} & =x\\

\sum_{j\in J} & f(a^{(j)})w^{(j)}=z.

\end{align*}

There's still one more wrinkle with which to contend. Other than possibly at extreme points, there will be more than one convex combination of mesh points producing the same $x$ vector, and they will not all produce the same approximation $z$ of $f(x)$. Consider an example in which $n=2$, $R=[2,5]\times[1,3]$, and $f(x)=x^{\prime}x=\left\Vert x\right\Vert ^{2}$. I'll use integer-valued mesh points. Assume that the optimal solution requires that $x=(3.5,2.2)$, in which case $f(x)=17.09$. Figure 1 illustrates the situation, with the values of $f()$ at the mesh points shown in red. (Click any figure to get a better resolution version in a separate browser tab.)

Figure 1: Rectangle $R$ and solution $x$ |

Figure 2: Weights that minimize $z$ |

Figure 3: Alternative weights that minimize $z$ |

Now assume instead that the solver prefers larger values of $f(x)$. The solution that maximizes $z$ is shown in Figure 4. It uses three corners of R, none of which are adjacent to $x,$ with weights 0.5, 0.4 and 0.1 Although this produces the correct value of $x$, the interpolated value of $z$ is $20.3$, which grossly overstates the correct value $17.09$.

Figure 4: Weights that maximize $z$ |

\[

\sum_{j=1}^{N_{i}}v_{i}^{(j)}=1\quad i=1,\dots,n.

\]Now we just need to add constraints forcing $w^{(j)}=0$ unless mesh point $j$ is one of the corners of the chosen hyperrectangle. Observe that, along any dimension $i$, the $i$-th component of any corner of the correct hyperrectangle will either be the chosen grid value or the next consecutive grid value. For instance, the rectangle containing $x$ in Figure 1 has lower left corner $(a_{1}^{(2)},a_{2}^{(2)})=(3,2)$. That means $v_{1}^{(2)}=1=v_{2}^{(2)}$. The corners of the rectangle have either $3=a_{1}^{(2)}$ or $4=a_{1}^{(3)}$ for their $x_{1}$ coordinate and either $2=a_{2}^{(2)}$ or $3=a_{2}^{(3)}$ for their $x_{2}$ coordinate.

So for $w^{(j)}$ to be nonzero, we need either $v_{i}^{(j_{i})}=1$ or $v_{i}^{(j_{i}-1)}=1$. This leads us to the constraints \[

w^{(j)}\le v_{i}^{(j_{i})}+v_{i}^{(j_{i-1})}

\]for all indices $j$ of mesh points and for all $i=1,\dots,n$, with the understanding that $v_{i}^{(0)}=0$ (since there is no grid point prior to the first one in any direction).

With those extra constraints, the solutions to our little example when we want $z$ small are unchanged, since they already obey the additional constraints. When we want $z$ large, however, the solution in Figure 4 is now infeasible. All three positive weights violate the new constraints. For instance, $w^{(1,3)}=0.5$ (the weight applied to the mesh point formed by the first grid value of $x_{1}$ and the third grid value of $x_{2}$), but $v_{1}^{(1)}+v_{1}^{(0)}=0$. The solutions that maximize $z$ end up being the same ones that minimize $z$ (those shown in Figures 2 and 3).

What remains is to take stock of how much the model has inflated. Let $$P=N_{1}\times\cdots\times N_{n}$$and$$S=N_{1}+\cdots+N_{n}.$$ We first added $P$ continuous variables ($w$) and $n+2$ constraints involving them. Then we added $S$ binary variables ($v$), $n$ constraints involving just them, and $nP$ constraints tying them to the $w$ variables. That's a grand total of $P$ continuous variables, $S$ binary variables and $2n+2+nP$ constraints. Note that $n\ll S\ll P<nP$. Also note that, in general, continuous variables are computationally cheaper than integer variables. As for the gaggle of extra constraints, modern solvers have ways to deal with some constraints "lazily", which mitigates the load to some extent.

An alternative approach would be to partion $R$ into nonoverlapping polygons (not necessarily hyperrectangles), assign weight variables $w$ to the corners of those polygons as above, and assign a binary variable to each polygon indicating whether it was selected. That would increase the number of binary variables from $S$ to $P$ (where $P$ would be the number of polygons) and decrease the number of added constraints from $2n+2+nP$ to $n+3+P.$ (The $n$ constraints requiring sums of binary variables to be 1 becomes a single constraint summing all the new binary variables. The $nP$ constraints tying the $w$ and $v$ variables together become $P$ constraints, one for each $w$ variable.) This approach is more flexible in terms of concentrating polygons where the curvature of $f$ is greatest, and it significantly reduces the number of constraints; but it significantly increases the number of binary variables, albeit with all of them tied into a single type 1 special ordered set. So it's hard for me to say which is better in general.

## Thursday, January 18, 2018

### More on "Core Points"

A few additions to yesterday's post occurred to me belatedly.

First, it may be a good idea to check whether your alleged core point $y^0$ is actually in the relative interior of the integer hull $\mathrm{conv}(Y)$. A sufficient condition is that, when you substitute $y^0$ into the constraints, all inequality constraints

Second, while the presence of equality constraints will definitely make the feasible region less than full dimension, it can occur even in a problem with only inequality constraints. Consider a problem with nonnegative general integer variables and the following constraints (as well as others, and other variables): \begin{align*} y_{1}+y_{2} +y_{3} & \le 2\\ y_{2}+y_{4} + y_{5} & \le 4\\ y_{1}+y_{3} + y_{4} + y_{5}& \ge 6 \end{align*} Although all the constraints are inequalities, the feasible region will live in the hyperplane $y_2 = 0$, and thus be less than full dimension. This points to why I said the condition in the previous paragraph is sufficient rather than necessary and sufficient for $y^0$ to be in the relative interior of the integer hull. In this example, there is no way to get positive slack in all three of the constraints (or, in fact, in any one of them) without violating the nonnegativity restriction on $y_2$.

Yesterday, I listed a few things one could try in the hope of getting a core point $y^0$ in the relative interior of the integer hull. Here are a few others that occurred to me. (Warning: I'm going to use the term "slack variable" for both slacks and surpluses.)

First, it may be a good idea to check whether your alleged core point $y^0$ is actually in the relative interior of the integer hull $\mathrm{conv}(Y)$. A sufficient condition is that, when you substitute $y^0$ into the constraints, all inequality constraints

*including variable bounds*have positive slack. (Equality constraints obviously will not have slack.) In particular, do not forget that nonnegativity restrictions count as bounds. If you specify $0\le y_i \le u_i$, then you are looking for $\epsilon \le y^0_i \le u_i - \epsilon$ for some $\epsilon > 0$ (and $\epsilon$ greater than your tolerance for rounding error).Second, while the presence of equality constraints will definitely make the feasible region less than full dimension, it can occur even in a problem with only inequality constraints. Consider a problem with nonnegative general integer variables and the following constraints (as well as others, and other variables): \begin{align*} y_{1}+y_{2} +y_{3} & \le 2\\ y_{2}+y_{4} + y_{5} & \le 4\\ y_{1}+y_{3} + y_{4} + y_{5}& \ge 6 \end{align*} Although all the constraints are inequalities, the feasible region will live in the hyperplane $y_2 = 0$, and thus be less than full dimension. This points to why I said the condition in the previous paragraph is sufficient rather than necessary and sufficient for $y^0$ to be in the relative interior of the integer hull. In this example, there is no way to get positive slack in all three of the constraints (or, in fact, in any one of them) without violating the nonnegativity restriction on $y_2$.

Yesterday, I listed a few things one could try in the hope of getting a core point $y^0$ in the relative interior of the integer hull. Here are a few others that occurred to me. (Warning: I'm going to use the term "slack variable" for both slacks and surpluses.)

- Tighten all inequality constraints (including variable bounds) and solve the LP relaxation of the tightened problem. (Feel free to change the objective function if you wish.) If you find a feasible solution $y^0$, it will be in the relative interior of the LP hull, and quite possibly in the integer hull. Good news: It's easy to do. Bad news: Even a modest tightening might make the problem infeasible (see example above).
- Insert slack variables in all constraints, including variable bounds. That means $0 \le y_i \le u_i$ would become \begin{align*} y_{i}-s_{i} & \ge0\\ y_{i}+t_{i} & \le u_{i} \end{align*} where $s_i \ge 0$ and $t_i \ge 0$ are slack variables. Maximize the minimum value of the slacks over the LP relaxation of the modified problem. Good news: If the solution has positive objective value (meaning all slacks are positive), you have a relative interior point of at least the LP hull. Bad news: An unfortunate combination of inequalities, like the example above, may prevent you from getting all slacks positive.

Labels:
Benders decomposition,
integer programming

## Wednesday, January 17, 2018

### Finding a "Core Point"

In a famous (or at least relatively famous) paper [1], Magnanti and Wong suggest a method to accelerate the progress of Benders decomposition for certain mixed-integer programs by sharpening "optimality" cuts. Their approach requires the determination of what they call a

I'll assume that any reader who made it this far knows what a convex hull is. That said, I should point out that the convex hull of $Y$ is

Mathematicians will know the term "relative interior", but it may not be familiar to OR/IE/MS people trying to solve MIP models. In Euclidean spaces, a point belongs to the interior of a set if there's a ball centered around that point and contained entirely in the set. In lay terms, you're at an interior point if you move slightly in any direction and not leave the set. In our context, that set is the integer hull, a convex polyhedron. The catch is that if the set is not full dimensional, there is no interior (or, more properly, the interior is empty). Picture the unit cube in $\mathbb{R}^3$ as the integer hull. The point $y^0 = (0.5, 0.5, 0.5)$ is in the interior; you can move a bit in any direction and stay in the cube. Now add the constraint $y_3 = 0.5$, which flattens the integer hull to a square (still containing $y^0$). You can move a bit in the $y_1$ and $y_2$ directions and stay in the square, but any vertical movement (parallel to the $y_3$ axis) and you're no longer feasible. That brings us to the relative interior. A point is in the relative interior of a set if it can be surrounded by a ball in the largest subspace for which the set is full dimensional, with the ball staying in the set. Returning to our example, before adding the equality constraint I could surround $y^0$ with a sphere contained in the unit cube. After adding the equality constraint, making the feasible region a square, I can no longer surround $y^0$ with a feasible sphere, but I can surround it with a feasible circle in the $y_3 = 0.5$ plane. So $y^0$ is at least in the relative interior of the integer hull after adding the equation constraint.

Back to Magnanti-Wong. To use their technique, you need to know a "core point" up front. In their paper, they mention that in some cases a core point will be obvious ... but in many cases it will not be. So I'll mention some techniques that

[1] Magnanti, T. L. & Wong, R. T., Accelerating Benders Decomposition: Algorithmic Enhancement and Model Selection Criteria.

*core point*. I'll try to follow their notation as much as possible. Let $Y$ be the set of integer-valued vectors satisfying all the constraints that involve solely the integer variables. Basically, $Y$ would be the feasible region for the problem if the continuous variables did not exist. A core point is a point $y^0$ in the relative interior of the convex hull of $Y$.I'll assume that any reader who made it this far knows what a convex hull is. That said, I should point out that the convex hull of $Y$ is

*not*the feasible region of the LP relaxation of $Y$. The feasible region of the LP relaxation is frequently termed the "LP hull", whereas the convex hull of $Y$ is typically referred to as the "integer hull". The integer hull is a subset (typically a proper subset) of the LP hull. Knowing what the integer hull is basically makes a MIP model pretty easy: solving a linear program over the integer hull yields the integer optimum. Since it's pretty well known that most MIP models are not easy, one can infer that in most cases the integer hull is not known (and not easy to figure out).Mathematicians will know the term "relative interior", but it may not be familiar to OR/IE/MS people trying to solve MIP models. In Euclidean spaces, a point belongs to the interior of a set if there's a ball centered around that point and contained entirely in the set. In lay terms, you're at an interior point if you move slightly in any direction and not leave the set. In our context, that set is the integer hull, a convex polyhedron. The catch is that if the set is not full dimensional, there is no interior (or, more properly, the interior is empty). Picture the unit cube in $\mathbb{R}^3$ as the integer hull. The point $y^0 = (0.5, 0.5, 0.5)$ is in the interior; you can move a bit in any direction and stay in the cube. Now add the constraint $y_3 = 0.5$, which flattens the integer hull to a square (still containing $y^0$). You can move a bit in the $y_1$ and $y_2$ directions and stay in the square, but any vertical movement (parallel to the $y_3$ axis) and you're no longer feasible. That brings us to the relative interior. A point is in the relative interior of a set if it can be surrounded by a ball in the largest subspace for which the set is full dimensional, with the ball staying in the set. Returning to our example, before adding the equality constraint I could surround $y^0$ with a sphere contained in the unit cube. After adding the equality constraint, making the feasible region a square, I can no longer surround $y^0$ with a feasible sphere, but I can surround it with a feasible circle in the $y_3 = 0.5$ plane. So $y^0$ is at least in the relative interior of the integer hull after adding the equation constraint.

Back to Magnanti-Wong. To use their technique, you need to know a "core point" up front. In their paper, they mention that in some cases a core point will be obvious ... but in many cases it will not be. So I'll mention some techniques that

*probably*yield a core point (but no general guarantee).- If you happen to know two or more integer-feasible points, average them. The average will be feasible, and unless all those points live on the same facet of the integer hull, you should get a relative interior point.
- Pick an objective function and optimize it (maximize or minimize, it doesn't matter) over $Y$, ignoring the continuous variables and any constraints involving them from the original problem. Do this a few times: minimize and maximize the same objective, switch the objective, iterate; or just minimize (or maximize) a different objective each time. I might use randomly generated objective functions, but an easy alternative is to minimize $y_1$, then maximize $y_1$, then minimize $y_2$, etc. Average the solutions you get. This is guaranteed to belong to integer hull, and almost surely (at least with random objectives and multiple iterations) to the relative interior of the integer hull. Bad news: you just solved a bunch of integer programs, which might be a trifle time-consuming.
- Use the previous method, but optimize over the LP hull (relaxing the integrality constraints on $y$). Again, average the solutions you get. Good news: Solving a handful of LPs is typically much less painful than solving a handful of IPs. Bad news: Since the LP hull is a superset of the integer hull, and since your LP solutions are all vertices of the LP hull, there's at least a chance that the average of them lives inside the LP hull but outside the integer hull. That said, I've used this method once or twice and not lost any sleep. If your core point happens to fall outside the integer hull, I don't think it will cause any problems; it just probably won't make your Benders decomposition solve any faster.
- Generate a bunch of random integer vectors of the correct dimension, and filter out those that do not satisfy the constraints defining $Y$. Average the survivors. Good news: Each of the surviving integer vectors will belong to the integer hull, so your averaged core point $y^0$ definitely will as well. Also, generating random integer vectors is pretty easy. Bad news: If the integer hull is less than full dimension, the probability of a random vector falling in it is zero, so the likelihood of any "survivors" is negligible. Mitigating news: In some cases you can randomly generate an integer vector and then tweak it to ensure it satisfies the constraints that are keeping the integer hull from being full dimension. For instance, suppose that you have a constraint of the form $\sum_i y_i = B$ where $B$ is an integer constant. Generate a random integer vector $\tilde{y}$, replace $\tilde{y}_1$ with $B-\sum_{i>1}\tilde{y}_i$, and you have a vector satisfying that constraint. If you can pull off similar tweaks for other equation constraints (staying integer-valued and not violating bounds on the variables), maybe you can get lucky and find a few integer-feasible points. In fact, even if you can only find one survivor (before exhausting your patience), you might get lucky and have it be in the relative interior of the integer hull. (You won't know this, but you can try using it as your core point and hope for the best.)

[1] Magnanti, T. L. & Wong, R. T., Accelerating Benders Decomposition: Algorithmic Enhancement and Model Selection Criteria.

*Operations Research*,**1981**,*29*, 464-484
Labels:
Benders decomposition,
integer programming

## Monday, January 1, 2018

### Ordering Index Vector with Java Streams

I bumped up against the following problem while doing some coding in Java 8 (and using streams where possible). Given a vector of objects $x_1, \dots, x_N$ that come from some domain having an ordering $\preccurlyeq$, find the vector of indices $i_1, \dots, i_N$ that sorts the original values into ascending order, i.e., such that $x_{i_1} \preccurlyeq x_{i_2} \preccurlyeq \cdots \preccurlyeq x_{i_N}$ I'm not sure there's an "official" name for that vector of indices, but I've seen it referred to more than once as the "ordering index vector" (hence the name of this post).

One of the nice features of the Java stream package is that it is really easy to sort streams, either using their natural ordering (where applicable) or using a specified ordering. As best I can tell, though, there is (at least currently) no built-in way to find out the original indices or positions of the sorted results. Fortunately, it didn't take to long to hack something that works. It may not be the most elegant way to get the ordering vector, but I'll share it anyway in case someone finds it useful.

My example will sort a vector of doubles, since that was what I was trying to do when I was forced to come up with this code. With fairly obvious modifications, it should work for sorting vectors of other types. Here is the code. Please try not to laugh.

The stream has to take a winding trip through the Ugly Forest to get this to work. We start out with an IntStream, because that is the easiest way to get a stream of indices. Unfortunately, sorting using a specified comparator is not supported by IntStream, so we have to "box" it get a stream of integers (Stream<Integer>). (Yes, fans, IntStream and stream of integer are two separate things.) The stream of integers is sorted by comparing the values they index in the original vector of double precision reals. The we use the identity function to map the stream of integers back to an IntStream (!) so that we can easily convert it to a vector of integers (meaning int[], not Integer[]).

When I said this might not be the "most elegant" approach, what I meant was that it looks clunky (at least to me). I look forward to being schooled in the comments section.

One of the nice features of the Java stream package is that it is really easy to sort streams, either using their natural ordering (where applicable) or using a specified ordering. As best I can tell, though, there is (at least currently) no built-in way to find out the original indices or positions of the sorted results. Fortunately, it didn't take to long to hack something that works. It may not be the most elegant way to get the ordering vector, but I'll share it anyway in case someone finds it useful.

My example will sort a vector of doubles, since that was what I was trying to do when I was forced to come up with this code. With fairly obvious modifications, it should work for sorting vectors of other types. Here is the code. Please try not to laugh.

// Create a vector of values whose order is desired. double[] vals = ... // Get the sort order for the values. int[] order = IntStream.range(0, vals.length) .boxed() .sorted((i, j) -> Double.compare(vals[i], vals[j])) .mapToInt(x -> x) .toArray(); // The sorted list is vals[order[0]], vals[order[1]], ...

The stream has to take a winding trip through the Ugly Forest to get this to work. We start out with an IntStream, because that is the easiest way to get a stream of indices. Unfortunately, sorting using a specified comparator is not supported by IntStream, so we have to "box" it get a stream of integers (Stream<Integer>). (Yes, fans, IntStream and stream of integer are two separate things.) The stream of integers is sorted by comparing the values they index in the original vector of double precision reals. The we use the identity function to map the stream of integers back to an IntStream (!) so that we can easily convert it to a vector of integers (meaning int[], not Integer[]).

When I said this might not be the "most elegant" approach, what I meant was that it looks clunky (at least to me). I look forward to being schooled in the comments section.

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